Adding upstream version 4.19.249.upstream/4.19.249

Signed-off-by: Daniel Baumann <daniel.baumann@progress-linux.org>

134 files changed, 35135 insertions, 0 deletions

diff --git a/Documentation/filesystems/00-INDEX b/Documentation/filesystems/00-INDEX
new file mode 100644
index 000000000..0937bade1
--- /dev/null
+++ b/Documentation/filesystems/00-INDEX
@@ -0,0 +1,153 @@
+00-INDEX
+	- this file (info on some of the filesystems supported by linux).
+Locking
+	- info on locking rules as they pertain to Linux VFS.
+9p.txt
+	- 9p (v9fs) is an implementation of the Plan 9 remote fs protocol.
+adfs.txt
+	- info and mount options for the Acorn Advanced Disc Filing System.
+afs.txt
+	- info and examples for the distributed AFS (Andrew File System) fs.
+affs.txt
+	- info and mount options for the Amiga Fast File System.
+autofs-mount-control.txt
+	- info on device control operations for autofs module.
+automount-support.txt
+	- information about filesystem automount support.
+befs.txt
+	- information about the BeOS filesystem for Linux.
+bfs.txt
+	- info for the SCO UnixWare Boot Filesystem (BFS).
+btrfs.txt
+	- info for the BTRFS filesystem.
+caching/
+	- directory containing filesystem cache documentation.
+ceph.txt
+	- info for the Ceph Distributed File System.
+cifs/
+	- directory containing CIFS filesystem documentation and example code.
+coda.txt
+	- description of the CODA filesystem.
+configfs/
+	- directory containing configfs documentation and example code.
+cramfs.txt
+	- info on the cram filesystem for small storage (ROMs etc).
+dax.txt
+	- info on avoiding the page cache for files stored on CPU-addressable
+	  storage devices.
+debugfs.txt
+	- info on the debugfs filesystem.
+devpts.txt
+	- info on the devpts filesystem.
+directory-locking
+	- info about the locking scheme used for directory operations.
+dlmfs.txt
+	- info on the userspace interface to the OCFS2 DLM.
+dnotify.txt
+	- info about directory notification in Linux.
+dnotify_test.c
+	- example program for dnotify.
+ecryptfs.txt
+	- docs on eCryptfs: stacked cryptographic filesystem for Linux.
+efivarfs.txt
+	- info for the efivarfs filesystem.
+exofs.txt
+	- info, usage, mount options, design about EXOFS.
+ext2.txt
+	- info, mount options and specifications for the Ext2 filesystem.
+ext3.txt
+	- info, mount options and specifications for the Ext3 filesystem.
+ext4.txt
+	- info, mount options and specifications for the Ext4 filesystem.
+f2fs.txt
+	- info and mount options for the F2FS filesystem.
+fiemap.txt
+	- info on fiemap ioctl.
+files.txt
+	- info on file management in the Linux kernel.
+fuse.txt
+	- info on the Filesystem in User SpacE including mount options.
+gfs2-glocks.txt
+	- info on the Global File System 2 - Glock internal locking rules.
+gfs2-uevents.txt
+	- info on the Global File System 2 - uevents.
+gfs2.txt
+	- info on the Global File System 2.
+hfs.txt
+	- info on the Macintosh HFS Filesystem for Linux.
+hfsplus.txt
+	- info on the Macintosh HFSPlus Filesystem for Linux.
+hpfs.txt
+	- info and mount options for the OS/2 HPFS.
+inotify.txt
+	- info on the powerful yet simple file change notification system.
+isofs.txt
+	- info and mount options for the ISO 9660 (CDROM) filesystem.
+jfs.txt
+	- info and mount options for the JFS filesystem.
+locks.txt
+	- info on file locking implementations, flock() vs. fcntl(), etc.
+mandatory-locking.txt
+	- info on the Linux implementation of Sys V mandatory file locking.
+nfs/
+	- nfs-related documentation.
+nilfs2.txt
+	- info and mount options for the NILFS2 filesystem.
+ntfs.txt
+	- info and mount options for the NTFS filesystem (Windows NT).
+ocfs2.txt
+	- info and mount options for the OCFS2 clustered filesystem.
+omfs.txt
+	- info on the Optimized MPEG FileSystem.
+path-lookup.txt
+	- info on path walking and name lookup locking.
+pohmelfs/
+	- directory containing pohmelfs filesystem documentation.
+porting
+	- various information on filesystem porting.
+proc.txt
+	- info on Linux's /proc filesystem.
+qnx6.txt
+	- info on the QNX6 filesystem.
+quota.txt
+	- info on Quota subsystem.
+ramfs-rootfs-initramfs.txt
+	- info on the 'in memory' filesystems ramfs, rootfs and initramfs.
+relay.txt
+	- info on relay, for efficient streaming from kernel to user space.
+romfs.txt
+	- description of the ROMFS filesystem.
+seq_file.txt
+	- how to use the seq_file API.
+sharedsubtree.txt
+	- a description of shared subtrees for namespaces.
+spufs.txt
+	- info and mount options for the SPU filesystem used on Cell.
+squashfs.txt
+	- info on the squashfs filesystem.
+sysfs-pci.txt
+	- info on accessing PCI device resources through sysfs.
+sysfs-tagging.txt
+	- info on sysfs tagging to avoid duplicates.
+sysfs.txt
+	- info on sysfs, a ram-based filesystem for exporting kernel objects.
+sysv-fs.txt
+	- info on the SystemV/V7/Xenix/Coherent filesystem.
+tmpfs.txt
+	- info on tmpfs, a filesystem that holds all files in virtual memory.
+ubifs.txt
+	- info on the Unsorted Block Images FileSystem.
+udf.txt
+	- info and mount options for the UDF filesystem.
+ufs.txt
+	- info on the ufs filesystem.
+vfat.txt
+	- info on using the VFAT filesystem used in Windows NT and Windows 95.
+vfs.txt
+	- overview of the Virtual File System.
+xfs-delayed-logging-design.txt
+	- info on the XFS Delayed Logging Design.
+xfs-self-describing-metadata.txt
+	- info on XFS Self Describing Metadata.
+xfs.txt
+	- info and mount options for the XFS filesystem.
diff --git a/Documentation/filesystems/9p.txt b/Documentation/filesystems/9p.txt
new file mode 100644
index 000000000..fec7144e8
--- /dev/null
+++ b/Documentation/filesystems/9p.txt
@@ -0,0 +1,161 @@
+	  	    v9fs: Plan 9 Resource Sharing for Linux
+		    =======================================
+
+ABOUT
+=====
+
+v9fs is a Unix implementation of the Plan 9 9p remote filesystem protocol.
+
+This software was originally developed by Ron Minnich <rminnich@sandia.gov>
+and Maya Gokhale.  Additional development by Greg Watson
+<gwatson@lanl.gov> and most recently Eric Van Hensbergen
+<ericvh@gmail.com>, Latchesar Ionkov <lucho@ionkov.net> and Russ Cox
+<rsc@swtch.com>.
+
+The best detailed explanation of the Linux implementation and applications of
+the 9p client is available in the form of a USENIX paper:
+   http://www.usenix.org/events/usenix05/tech/freenix/hensbergen.html
+
+Other applications are described in the following papers:
+	* XCPU & Clustering
+		http://xcpu.org/papers/xcpu-talk.pdf
+	* KVMFS: control file system for KVM
+		http://xcpu.org/papers/kvmfs.pdf
+	* CellFS: A New Programming Model for the Cell BE
+		http://xcpu.org/papers/cellfs-talk.pdf
+	* PROSE I/O: Using 9p to enable Application Partitions
+		http://plan9.escet.urjc.es/iwp9/cready/PROSE_iwp9_2006.pdf
+	* VirtFS: A Virtualization Aware File System pass-through
+		http://goo.gl/3WPDg
+
+USAGE
+=====
+
+For remote file server:
+
+	mount -t 9p 10.10.1.2 /mnt/9
+
+For Plan 9 From User Space applications (http://swtch.com/plan9)
+
+	mount -t 9p `namespace`/acme /mnt/9 -o trans=unix,uname=$USER
+
+For server running on QEMU host with virtio transport:
+
+	mount -t 9p -o trans=virtio <mount_tag> /mnt/9
+
+where mount_tag is the tag associated by the server to each of the exported
+mount points. Each 9P export is seen by the client as a virtio device with an
+associated "mount_tag" property. Available mount tags can be
+seen by reading /sys/bus/virtio/drivers/9pnet_virtio/virtio<n>/mount_tag files.
+
+OPTIONS
+=======
+
+  trans=name	select an alternative transport.  Valid options are
+  		currently:
+			unix 	- specifying a named pipe mount point
+			tcp	- specifying a normal TCP/IP connection
+			fd   	- used passed file descriptors for connection
+                                (see rfdno and wfdno)
+			virtio	- connect to the next virtio channel available
+				(from QEMU with trans_virtio module)
+			rdma	- connect to a specified RDMA channel
+
+  uname=name	user name to attempt mount as on the remote server.  The
+  		server may override or ignore this value.  Certain user
+		names may require authentication.
+
+  aname=name	aname specifies the file tree to access when the server is
+  		offering several exported file systems.
+
+  cache=mode	specifies a caching policy.  By default, no caches are used.
+                        none = default no cache policy, metadata and data
+                                alike are synchronous.
+			loose = no attempts are made at consistency,
+                                intended for exclusive, read-only mounts
+                        fscache = use FS-Cache for a persistent, read-only
+				cache backend.
+                        mmap = minimal cache that is only used for read-write
+                                mmap.  Northing else is cached, like cache=none
+
+  debug=n	specifies debug level.  The debug level is a bitmask.
+			0x01  = display verbose error messages
+			0x02  = developer debug (DEBUG_CURRENT)
+			0x04  = display 9p trace
+			0x08  = display VFS trace
+			0x10  = display Marshalling debug
+			0x20  = display RPC debug
+			0x40  = display transport debug
+			0x80  = display allocation debug
+			0x100 = display protocol message debug
+			0x200 = display Fid debug
+			0x400 = display packet debug
+			0x800 = display fscache tracing debug
+
+  rfdno=n	the file descriptor for reading with trans=fd
+
+  wfdno=n	the file descriptor for writing with trans=fd
+
+  msize=n	the number of bytes to use for 9p packet payload
+
+  port=n	port to connect to on the remote server
+
+  noextend	force legacy mode (no 9p2000.u or 9p2000.L semantics)
+
+  version=name	Select 9P protocol version. Valid options are:
+			9p2000          - Legacy mode (same as noextend)
+			9p2000.u        - Use 9P2000.u protocol
+			9p2000.L        - Use 9P2000.L protocol
+
+  dfltuid	attempt to mount as a particular uid
+
+  dfltgid	attempt to mount with a particular gid
+
+  afid		security channel - used by Plan 9 authentication protocols
+
+  nodevmap	do not map special files - represent them as normal files.
+  		This can be used to share devices/named pipes/sockets between
+		hosts.  This functionality will be expanded in later versions.
+
+  access	there are four access modes.
+			user  = if a user tries to access a file on v9fs
+			        filesystem for the first time, v9fs sends an
+			        attach command (Tattach) for that user.
+				This is the default mode.
+			<uid> = allows only user with uid=<uid> to access
+				the files on the mounted filesystem
+			any   = v9fs does single attach and performs all
+				operations as one user
+			client = ACL based access check on the 9p client
+			         side for access validation
+
+  cachetag	cache tag to use the specified persistent cache.
+		cache tags for existing cache sessions can be listed at
+		/sys/fs/9p/caches. (applies only to cache=fscache)
+
+RESOURCES
+=========
+
+Protocol specifications are maintained on github:
+http://ericvh.github.com/9p-rfc/
+
+9p client and server implementations are listed on
+http://9p.cat-v.org/implementations
+
+A 9p2000.L server is being developed by LLNL and can be found
+at http://code.google.com/p/diod/
+
+There are user and developer mailing lists available through the v9fs project
+on sourceforge (http://sourceforge.net/projects/v9fs).
+
+News and other information is maintained on a Wiki.
+(http://sf.net/apps/mediawiki/v9fs/index.php).
+
+Bug reports are best issued via the mailing list.
+
+For more information on the Plan 9 Operating System check out
+http://plan9.bell-labs.com/plan9
+
+For information on Plan 9 from User Space (Plan 9 applications and libraries
+ported to Linux/BSD/OSX/etc) check out http://swtch.com/plan9
+
diff --git a/Documentation/filesystems/Locking b/Documentation/filesystems/Locking
new file mode 100644
index 000000000..efea228cc
--- /dev/null
+++ b/Documentation/filesystems/Locking
@@ -0,0 +1,588 @@
+	The text below describes the locking rules for VFS-related methods.
+It is (believed to be) up-to-date. *Please*, if you change anything in
+prototypes or locking protocols - update this file. And update the relevant
+instances in the tree, don't leave that to maintainers of filesystems/devices/
+etc. At the very least, put the list of dubious cases in the end of this file.
+Don't turn it into log - maintainers of out-of-the-tree code are supposed to
+be able to use diff(1).
+	Thing currently missing here: socket operations. Alexey?
+
+--------------------------- dentry_operations --------------------------
+prototypes:
+	int (*d_revalidate)(struct dentry *, unsigned int);
+	int (*d_weak_revalidate)(struct dentry *, unsigned int);
+	int (*d_hash)(const struct dentry *, struct qstr *);
+	int (*d_compare)(const struct dentry *,
+			unsigned int, const char *, const struct qstr *);
+	int (*d_delete)(struct dentry *);
+	int (*d_init)(struct dentry *);
+	void (*d_release)(struct dentry *);
+	void (*d_iput)(struct dentry *, struct inode *);
+	char *(*d_dname)((struct dentry *dentry, char *buffer, int buflen);
+	struct vfsmount *(*d_automount)(struct path *path);
+	int (*d_manage)(const struct path *, bool);
+	struct dentry *(*d_real)(struct dentry *, const struct inode *);
+
+locking rules:
+		rename_lock	->d_lock	may block	rcu-walk
+d_revalidate:	no		no		yes (ref-walk)	maybe
+d_weak_revalidate:no		no		yes	 	no
+d_hash		no		no		no		maybe
+d_compare:	yes		no		no		maybe
+d_delete:	no		yes		no		no
+d_init:	no		no		yes		no
+d_release:	no		no		yes		no
+d_prune:        no              yes             no              no
+d_iput:		no		no		yes		no
+d_dname:	no		no		no		no
+d_automount:	no		no		yes		no
+d_manage:	no		no		yes (ref-walk)	maybe
+d_real		no		no		yes 		no
+
+--------------------------- inode_operations --------------------------- 
+prototypes:
+	int (*create) (struct inode *,struct dentry *,umode_t, bool);
+	struct dentry * (*lookup) (struct inode *,struct dentry *, unsigned int);
+	int (*link) (struct dentry *,struct inode *,struct dentry *);
+	int (*unlink) (struct inode *,struct dentry *);
+	int (*symlink) (struct inode *,struct dentry *,const char *);
+	int (*mkdir) (struct inode *,struct dentry *,umode_t);
+	int (*rmdir) (struct inode *,struct dentry *);
+	int (*mknod) (struct inode *,struct dentry *,umode_t,dev_t);
+	int (*rename) (struct inode *, struct dentry *,
+			struct inode *, struct dentry *, unsigned int);
+	int (*readlink) (struct dentry *, char __user *,int);
+	const char *(*get_link) (struct dentry *, struct inode *, void **);
+	void (*truncate) (struct inode *);
+	int (*permission) (struct inode *, int, unsigned int);
+	int (*get_acl)(struct inode *, int);
+	int (*setattr) (struct dentry *, struct iattr *);
+	int (*getattr) (const struct path *, struct kstat *, u32, unsigned int);
+	ssize_t (*listxattr) (struct dentry *, char *, size_t);
+	int (*fiemap)(struct inode *, struct fiemap_extent_info *, u64 start, u64 len);
+	void (*update_time)(struct inode *, struct timespec *, int);
+	int (*atomic_open)(struct inode *, struct dentry *,
+				struct file *, unsigned open_flag,
+				umode_t create_mode);
+	int (*tmpfile) (struct inode *, struct dentry *, umode_t);
+
+locking rules:
+	all may block
+		i_rwsem(inode)
+lookup:		shared
+create:		exclusive
+link:		exclusive (both)
+mknod:		exclusive
+symlink:	exclusive
+mkdir:		exclusive
+unlink:		exclusive (both)
+rmdir:		exclusive (both)(see below)
+rename:		exclusive (all)	(see below)
+readlink:	no
+get_link:	no
+setattr:	exclusive
+permission:	no (may not block if called in rcu-walk mode)
+get_acl:	no
+getattr:	no
+listxattr:	no
+fiemap:		no
+update_time:	no
+atomic_open:	exclusive
+tmpfile:	no
+
+
+	Additionally, ->rmdir(), ->unlink() and ->rename() have ->i_rwsem
+	exclusive on victim.
+	cross-directory ->rename() has (per-superblock) ->s_vfs_rename_sem.
+
+See Documentation/filesystems/directory-locking for more detailed discussion
+of the locking scheme for directory operations.
+
+----------------------- xattr_handler operations -----------------------
+prototypes:
+	bool (*list)(struct dentry *dentry);
+	int (*get)(const struct xattr_handler *handler, struct dentry *dentry,
+		   struct inode *inode, const char *name, void *buffer,
+		   size_t size);
+	int (*set)(const struct xattr_handler *handler, struct dentry *dentry,
+		   struct inode *inode, const char *name, const void *buffer,
+		   size_t size, int flags);
+
+locking rules:
+	all may block
+		i_rwsem(inode)
+list:		no
+get:		no
+set:		exclusive
+
+--------------------------- super_operations ---------------------------
+prototypes:
+	struct inode *(*alloc_inode)(struct super_block *sb);
+	void (*destroy_inode)(struct inode *);
+	void (*dirty_inode) (struct inode *, int flags);
+	int (*write_inode) (struct inode *, struct writeback_control *wbc);
+	int (*drop_inode) (struct inode *);
+	void (*evict_inode) (struct inode *);
+	void (*put_super) (struct super_block *);
+	int (*sync_fs)(struct super_block *sb, int wait);
+	int (*freeze_fs) (struct super_block *);
+	int (*unfreeze_fs) (struct super_block *);
+	int (*statfs) (struct dentry *, struct kstatfs *);
+	int (*remount_fs) (struct super_block *, int *, char *);
+	void (*umount_begin) (struct super_block *);
+	int (*show_options)(struct seq_file *, struct dentry *);
+	ssize_t (*quota_read)(struct super_block *, int, char *, size_t, loff_t);
+	ssize_t (*quota_write)(struct super_block *, int, const char *, size_t, loff_t);
+	int (*bdev_try_to_free_page)(struct super_block*, struct page*, gfp_t);
+
+locking rules:
+	All may block [not true, see below]
+			s_umount
+alloc_inode:
+destroy_inode:
+dirty_inode:
+write_inode:
+drop_inode:				!!!inode->i_lock!!!
+evict_inode:
+put_super:		write
+sync_fs:		read
+freeze_fs:		write
+unfreeze_fs:		write
+statfs:			maybe(read)	(see below)
+remount_fs:		write
+umount_begin:		no
+show_options:		no		(namespace_sem)
+quota_read:		no		(see below)
+quota_write:		no		(see below)
+bdev_try_to_free_page:	no		(see below)
+
+->statfs() has s_umount (shared) when called by ustat(2) (native or
+compat), but that's an accident of bad API; s_umount is used to pin
+the superblock down when we only have dev_t given us by userland to
+identify the superblock.  Everything else (statfs(), fstatfs(), etc.)
+doesn't hold it when calling ->statfs() - superblock is pinned down
+by resolving the pathname passed to syscall.
+->quota_read() and ->quota_write() functions are both guaranteed to
+be the only ones operating on the quota file by the quota code (via
+dqio_sem) (unless an admin really wants to screw up something and
+writes to quota files with quotas on). For other details about locking
+see also dquot_operations section.
+->bdev_try_to_free_page is called from the ->releasepage handler of
+the block device inode.  See there for more details.
+
+--------------------------- file_system_type ---------------------------
+prototypes:
+	struct dentry *(*mount) (struct file_system_type *, int,
+		       const char *, void *);
+	void (*kill_sb) (struct super_block *);
+locking rules:
+		may block
+mount		yes
+kill_sb		yes
+
+->mount() returns ERR_PTR or the root dentry; its superblock should be locked
+on return.
+->kill_sb() takes a write-locked superblock, does all shutdown work on it,
+unlocks and drops the reference.
+
+--------------------------- address_space_operations --------------------------
+prototypes:
+	int (*writepage)(struct page *page, struct writeback_control *wbc);
+	int (*readpage)(struct file *, struct page *);
+	int (*writepages)(struct address_space *, struct writeback_control *);
+	int (*set_page_dirty)(struct page *page);
+	int (*readpages)(struct file *filp, struct address_space *mapping,
+			struct list_head *pages, unsigned nr_pages);
+	int (*write_begin)(struct file *, struct address_space *mapping,
+				loff_t pos, unsigned len, unsigned flags,
+				struct page **pagep, void **fsdata);
+	int (*write_end)(struct file *, struct address_space *mapping,
+				loff_t pos, unsigned len, unsigned copied,
+				struct page *page, void *fsdata);
+	sector_t (*bmap)(struct address_space *, sector_t);
+	void (*invalidatepage) (struct page *, unsigned int, unsigned int);
+	int (*releasepage) (struct page *, int);
+	void (*freepage)(struct page *);
+	int (*direct_IO)(struct kiocb *, struct iov_iter *iter);
+	bool (*isolate_page) (struct page *, isolate_mode_t);
+	int (*migratepage)(struct address_space *, struct page *, struct page *);
+	void (*putback_page) (struct page *);
+	int (*launder_page)(struct page *);
+	int (*is_partially_uptodate)(struct page *, unsigned long, unsigned long);
+	int (*error_remove_page)(struct address_space *, struct page *);
+	int (*swap_activate)(struct file *);
+	int (*swap_deactivate)(struct file *);
+
+locking rules:
+	All except set_page_dirty and freepage may block
+
+			PageLocked(page)	i_rwsem
+writepage:		yes, unlocks (see below)
+readpage:		yes, unlocks
+writepages:
+set_page_dirty		no
+readpages:
+write_begin:		locks the page		exclusive
+write_end:		yes, unlocks		exclusive
+bmap:
+invalidatepage:		yes
+releasepage:		yes
+freepage:		yes
+direct_IO:
+isolate_page:		yes
+migratepage:		yes (both)
+putback_page:		yes
+launder_page:		yes
+is_partially_uptodate:	yes
+error_remove_page:	yes
+swap_activate:		no
+swap_deactivate:	no
+
+	->write_begin(), ->write_end() and ->readpage() may be called from
+the request handler (/dev/loop).
+
+	->readpage() unlocks the page, either synchronously or via I/O
+completion.
+
+	->readpages() populates the pagecache with the passed pages and starts
+I/O against them.  They come unlocked upon I/O completion.
+
+	->writepage() is used for two purposes: for "memory cleansing" and for
+"sync".  These are quite different operations and the behaviour may differ
+depending upon the mode.
+
+If writepage is called for sync (wbc->sync_mode != WBC_SYNC_NONE) then
+it *must* start I/O against the page, even if that would involve
+blocking on in-progress I/O.
+
+If writepage is called for memory cleansing (sync_mode ==
+WBC_SYNC_NONE) then its role is to get as much writeout underway as
+possible.  So writepage should try to avoid blocking against
+currently-in-progress I/O.
+
+If the filesystem is not called for "sync" and it determines that it
+would need to block against in-progress I/O to be able to start new I/O
+against the page the filesystem should redirty the page with
+redirty_page_for_writepage(), then unlock the page and return zero.
+This may also be done to avoid internal deadlocks, but rarely.
+
+If the filesystem is called for sync then it must wait on any
+in-progress I/O and then start new I/O.
+
+The filesystem should unlock the page synchronously, before returning to the
+caller, unless ->writepage() returns special WRITEPAGE_ACTIVATE
+value. WRITEPAGE_ACTIVATE means that page cannot really be written out
+currently, and VM should stop calling ->writepage() on this page for some
+time. VM does this by moving page to the head of the active list, hence the
+name.
+
+Unless the filesystem is going to redirty_page_for_writepage(), unlock the page
+and return zero, writepage *must* run set_page_writeback() against the page,
+followed by unlocking it.  Once set_page_writeback() has been run against the
+page, write I/O can be submitted and the write I/O completion handler must run
+end_page_writeback() once the I/O is complete.  If no I/O is submitted, the
+filesystem must run end_page_writeback() against the page before returning from
+writepage.
+
+That is: after 2.5.12, pages which are under writeout are *not* locked.  Note,
+if the filesystem needs the page to be locked during writeout, that is ok, too,
+the page is allowed to be unlocked at any point in time between the calls to
+set_page_writeback() and end_page_writeback().
+
+Note, failure to run either redirty_page_for_writepage() or the combination of
+set_page_writeback()/end_page_writeback() on a page submitted to writepage
+will leave the page itself marked clean but it will be tagged as dirty in the
+radix tree.  This incoherency can lead to all sorts of hard-to-debug problems
+in the filesystem like having dirty inodes at umount and losing written data.
+
+	->writepages() is used for periodic writeback and for syscall-initiated
+sync operations.  The address_space should start I/O against at least
+*nr_to_write pages.  *nr_to_write must be decremented for each page which is
+written.  The address_space implementation may write more (or less) pages
+than *nr_to_write asks for, but it should try to be reasonably close.  If
+nr_to_write is NULL, all dirty pages must be written.
+
+writepages should _only_ write pages which are present on
+mapping->io_pages.
+
+	->set_page_dirty() is called from various places in the kernel
+when the target page is marked as needing writeback.  It may be called
+under spinlock (it cannot block) and is sometimes called with the page
+not locked.
+
+	->bmap() is currently used by legacy ioctl() (FIBMAP) provided by some
+filesystems and by the swapper. The latter will eventually go away.  Please,
+keep it that way and don't breed new callers.
+
+	->invalidatepage() is called when the filesystem must attempt to drop
+some or all of the buffers from the page when it is being truncated. It
+returns zero on success. If ->invalidatepage is zero, the kernel uses
+block_invalidatepage() instead.
+
+	->releasepage() is called when the kernel is about to try to drop the
+buffers from the page in preparation for freeing it.  It returns zero to
+indicate that the buffers are (or may be) freeable.  If ->releasepage is zero,
+the kernel assumes that the fs has no private interest in the buffers.
+
+	->freepage() is called when the kernel is done dropping the page
+from the page cache.
+
+	->launder_page() may be called prior to releasing a page if
+it is still found to be dirty. It returns zero if the page was successfully
+cleaned, or an error value if not. Note that in order to prevent the page
+getting mapped back in and redirtied, it needs to be kept locked
+across the entire operation.
+
+	->swap_activate will be called with a non-zero argument on
+files backing (non block device backed) swapfiles. A return value
+of zero indicates success, in which case this file can be used for
+backing swapspace. The swapspace operations will be proxied to the
+address space operations.
+
+	->swap_deactivate() will be called in the sys_swapoff()
+path after ->swap_activate() returned success.
+
+----------------------- file_lock_operations ------------------------------
+prototypes:
+	void (*fl_copy_lock)(struct file_lock *, struct file_lock *);
+	void (*fl_release_private)(struct file_lock *);
+
+
+locking rules:
+			inode->i_lock	may block
+fl_copy_lock:		yes		no
+fl_release_private:	maybe		maybe[1]
+
+[1]:	->fl_release_private for flock or POSIX locks is currently allowed
+to block. Leases however can still be freed while the i_lock is held and
+so fl_release_private called on a lease should not block.
+
+----------------------- lock_manager_operations ---------------------------
+prototypes:
+	int (*lm_compare_owner)(struct file_lock *, struct file_lock *);
+	unsigned long (*lm_owner_key)(struct file_lock *);
+	void (*lm_notify)(struct file_lock *);  /* unblock callback */
+	int (*lm_grant)(struct file_lock *, struct file_lock *, int);
+	void (*lm_break)(struct file_lock *); /* break_lease callback */
+	int (*lm_change)(struct file_lock **, int);
+
+locking rules:
+
+			inode->i_lock	blocked_lock_lock	may block
+lm_compare_owner:	yes[1]		maybe			no
+lm_owner_key		yes[1]		yes			no
+lm_notify:		yes		yes			no
+lm_grant:		no		no			no
+lm_break:		yes		no			no
+lm_change		yes		no			no
+
+[1]:	->lm_compare_owner and ->lm_owner_key are generally called with
+*an* inode->i_lock held. It may not be the i_lock of the inode
+associated with either file_lock argument! This is the case with deadlock
+detection, since the code has to chase down the owners of locks that may
+be entirely unrelated to the one on which the lock is being acquired.
+For deadlock detection however, the blocked_lock_lock is also held. The
+fact that these locks are held ensures that the file_locks do not
+disappear out from under you while doing the comparison or generating an
+owner key.
+
+--------------------------- buffer_head -----------------------------------
+prototypes:
+	void (*b_end_io)(struct buffer_head *bh, int uptodate);
+
+locking rules:
+	called from interrupts. In other words, extreme care is needed here.
+bh is locked, but that's all warranties we have here. Currently only RAID1,
+highmem, fs/buffer.c, and fs/ntfs/aops.c are providing these. Block devices
+call this method upon the IO completion.
+
+--------------------------- block_device_operations -----------------------
+prototypes:
+	int (*open) (struct block_device *, fmode_t);
+	int (*release) (struct gendisk *, fmode_t);
+	int (*ioctl) (struct block_device *, fmode_t, unsigned, unsigned long);
+	int (*compat_ioctl) (struct block_device *, fmode_t, unsigned, unsigned long);
+	int (*direct_access) (struct block_device *, sector_t, void **,
+				unsigned long *);
+	int (*media_changed) (struct gendisk *);
+	void (*unlock_native_capacity) (struct gendisk *);
+	int (*revalidate_disk) (struct gendisk *);
+	int (*getgeo)(struct block_device *, struct hd_geometry *);
+	void (*swap_slot_free_notify) (struct block_device *, unsigned long);
+
+locking rules:
+			bd_mutex
+open:			yes
+release:		yes
+ioctl:			no
+compat_ioctl:		no
+direct_access:		no
+media_changed:		no
+unlock_native_capacity:	no
+revalidate_disk:	no
+getgeo:			no
+swap_slot_free_notify:	no	(see below)
+
+media_changed, unlock_native_capacity and revalidate_disk are called only from
+check_disk_change().
+
+swap_slot_free_notify is called with swap_lock and sometimes the page lock
+held.
+
+
+--------------------------- file_operations -------------------------------
+prototypes:
+	loff_t (*llseek) (struct file *, loff_t, int);
+	ssize_t (*read) (struct file *, char __user *, size_t, loff_t *);
+	ssize_t (*write) (struct file *, const char __user *, size_t, loff_t *);
+	ssize_t (*read_iter) (struct kiocb *, struct iov_iter *);
+	ssize_t (*write_iter) (struct kiocb *, struct iov_iter *);
+	int (*iterate) (struct file *, struct dir_context *);
+	int (*iterate_shared) (struct file *, struct dir_context *);
+	__poll_t (*poll) (struct file *, struct poll_table_struct *);
+	long (*unlocked_ioctl) (struct file *, unsigned int, unsigned long);
+	long (*compat_ioctl) (struct file *, unsigned int, unsigned long);
+	int (*mmap) (struct file *, struct vm_area_struct *);
+	int (*open) (struct inode *, struct file *);
+	int (*flush) (struct file *);
+	int (*release) (struct inode *, struct file *);
+	int (*fsync) (struct file *, loff_t start, loff_t end, int datasync);
+	int (*fasync) (int, struct file *, int);
+	int (*lock) (struct file *, int, struct file_lock *);
+	ssize_t (*readv) (struct file *, const struct iovec *, unsigned long,
+			loff_t *);
+	ssize_t (*writev) (struct file *, const struct iovec *, unsigned long,
+			loff_t *);
+	ssize_t (*sendfile) (struct file *, loff_t *, size_t, read_actor_t,
+			void __user *);
+	ssize_t (*sendpage) (struct file *, struct page *, int, size_t,
+			loff_t *, int);
+	unsigned long (*get_unmapped_area)(struct file *, unsigned long,
+			unsigned long, unsigned long, unsigned long);
+	int (*check_flags)(int);
+	int (*flock) (struct file *, int, struct file_lock *);
+	ssize_t (*splice_write)(struct pipe_inode_info *, struct file *, loff_t *,
+			size_t, unsigned int);
+	ssize_t (*splice_read)(struct file *, loff_t *, struct pipe_inode_info *,
+			size_t, unsigned int);
+	int (*setlease)(struct file *, long, struct file_lock **, void **);
+	long (*fallocate)(struct file *, int, loff_t, loff_t);
+};
+
+locking rules:
+	All may block.
+
+->llseek() locking has moved from llseek to the individual llseek
+implementations.  If your fs is not using generic_file_llseek, you
+need to acquire and release the appropriate locks in your ->llseek().
+For many filesystems, it is probably safe to acquire the inode
+mutex or just to use i_size_read() instead.
+Note: this does not protect the file->f_pos against concurrent modifications
+since this is something the userspace has to take care about.
+
+->iterate() is called with i_rwsem exclusive.
+
+->iterate_shared() is called with i_rwsem at least shared.
+
+->fasync() is responsible for maintaining the FASYNC bit in filp->f_flags.
+Most instances call fasync_helper(), which does that maintenance, so it's
+not normally something one needs to worry about.  Return values > 0 will be
+mapped to zero in the VFS layer.
+
+->readdir() and ->ioctl() on directories must be changed. Ideally we would
+move ->readdir() to inode_operations and use a separate method for directory
+->ioctl() or kill the latter completely. One of the problems is that for
+anything that resembles union-mount we won't have a struct file for all
+components. And there are other reasons why the current interface is a mess...
+
+->read on directories probably must go away - we should just enforce -EISDIR
+in sys_read() and friends.
+
+->setlease operations should call generic_setlease() before or after setting
+the lease within the individual filesystem to record the result of the
+operation
+
+--------------------------- dquot_operations -------------------------------
+prototypes:
+	int (*write_dquot) (struct dquot *);
+	int (*acquire_dquot) (struct dquot *);
+	int (*release_dquot) (struct dquot *);
+	int (*mark_dirty) (struct dquot *);
+	int (*write_info) (struct super_block *, int);
+
+These operations are intended to be more or less wrapping functions that ensure
+a proper locking wrt the filesystem and call the generic quota operations.
+
+What filesystem should expect from the generic quota functions:
+
+		FS recursion	Held locks when called
+write_dquot:	yes		dqonoff_sem or dqptr_sem
+acquire_dquot:	yes		dqonoff_sem or dqptr_sem
+release_dquot:	yes		dqonoff_sem or dqptr_sem
+mark_dirty:	no		-
+write_info:	yes		dqonoff_sem
+
+FS recursion means calling ->quota_read() and ->quota_write() from superblock
+operations.
+
+More details about quota locking can be found in fs/dquot.c.
+
+--------------------------- vm_operations_struct -----------------------------
+prototypes:
+	void (*open)(struct vm_area_struct*);
+	void (*close)(struct vm_area_struct*);
+	vm_fault_t (*fault)(struct vm_area_struct*, struct vm_fault *);
+	vm_fault_t (*page_mkwrite)(struct vm_area_struct *, struct vm_fault *);
+	vm_fault_t (*pfn_mkwrite)(struct vm_area_struct *, struct vm_fault *);
+	int (*access)(struct vm_area_struct *, unsigned long, void*, int, int);
+
+locking rules:
+		mmap_sem	PageLocked(page)
+open:		yes
+close:		yes
+fault:		yes		can return with page locked
+map_pages:	yes
+page_mkwrite:	yes		can return with page locked
+pfn_mkwrite:	yes
+access:		yes
+
+	->fault() is called when a previously not present pte is about
+to be faulted in. The filesystem must find and return the page associated
+with the passed in "pgoff" in the vm_fault structure. If it is possible that
+the page may be truncated and/or invalidated, then the filesystem must lock
+the page, then ensure it is not already truncated (the page lock will block
+subsequent truncate), and then return with VM_FAULT_LOCKED, and the page
+locked. The VM will unlock the page.
+
+	->map_pages() is called when VM asks to map easy accessible pages.
+Filesystem should find and map pages associated with offsets from "start_pgoff"
+till "end_pgoff". ->map_pages() is called with page table locked and must
+not block.  If it's not possible to reach a page without blocking,
+filesystem should skip it. Filesystem should use do_set_pte() to setup
+page table entry. Pointer to entry associated with the page is passed in
+"pte" field in vm_fault structure. Pointers to entries for other offsets
+should be calculated relative to "pte".
+
+	->page_mkwrite() is called when a previously read-only pte is
+about to become writeable. The filesystem again must ensure that there are
+no truncate/invalidate races, and then return with the page locked. If
+the page has been truncated, the filesystem should not look up a new page
+like the ->fault() handler, but simply return with VM_FAULT_NOPAGE, which
+will cause the VM to retry the fault.
+
+	->pfn_mkwrite() is the same as page_mkwrite but when the pte is
+VM_PFNMAP or VM_MIXEDMAP with a page-less entry. Expected return is
+VM_FAULT_NOPAGE. Or one of the VM_FAULT_ERROR types. The default behavior
+after this call is to make the pte read-write, unless pfn_mkwrite returns
+an error.
+
+	->access() is called when get_user_pages() fails in
+access_process_vm(), typically used to debug a process through
+/proc/pid/mem or ptrace.  This function is needed only for
+VM_IO | VM_PFNMAP VMAs.
+
+================================================================================
+			Dubious stuff
+
+(if you break something or notice that it is broken and do not fix it yourself
+- at least put it here)
diff --git a/Documentation/filesystems/adfs.txt b/Documentation/filesystems/adfs.txt
new file mode 100644
index 000000000..594976635
--- /dev/null
+++ b/Documentation/filesystems/adfs.txt
@@ -0,0 +1,75 @@
+Mount options for ADFS
+----------------------
+
+  uid=nnn	All files in the partition will be owned by
+		user id nnn.  Default 0 (root).
+  gid=nnn	All files in the partition will be in group
+		nnn.  Default 0 (root).
+  ownmask=nnn	The permission mask for ADFS 'owner' permissions
+		will be nnn.  Default 0700.
+  othmask=nnn	The permission mask for ADFS 'other' permissions
+		will be nnn.  Default 0077.
+  ftsuffix=n	When ftsuffix=0, no file type suffix will be applied.
+		When ftsuffix=1, a hexadecimal suffix corresponding to
+		the RISC OS file type will be added.  Default 0.
+
+Mapping of ADFS permissions to Linux permissions
+------------------------------------------------
+
+  ADFS permissions consist of the following:
+
+	Owner read
+	Owner write
+	Other read
+	Other write
+
+  (In older versions, an 'execute' permission did exist, but this
+   does not hold the same meaning as the Linux 'execute' permission
+   and is now obsolete).
+
+  The mapping is performed as follows:
+
+	Owner read				-> -r--r--r--
+	Owner write				-> --w--w---w
+	Owner read and filetype UnixExec	-> ---x--x--x
+    These are then masked by ownmask, eg 700	-> -rwx------
+	Possible owner mode permissions		-> -rwx------
+
+	Other read				-> -r--r--r--
+	Other write				-> --w--w--w-
+	Other read and filetype UnixExec	-> ---x--x--x
+    These are then masked by othmask, eg 077	-> ----rwxrwx
+	Possible other mode permissions		-> ----rwxrwx
+
+  Hence, with the default masks, if a file is owner read/write, and
+  not a UnixExec filetype, then the permissions will be:
+
+			-rw-------
+
+  However, if the masks were ownmask=0770,othmask=0007, then this would
+  be modified to:
+			-rw-rw----
+
+  There is no restriction on what you can do with these masks.  You may
+  wish that either read bits give read access to the file for all, but
+  keep the default write protection (ownmask=0755,othmask=0577):
+
+			-rw-r--r--
+
+  You can therefore tailor the permission translation to whatever you
+  desire the permissions should be under Linux.
+
+RISC OS file type suffix
+------------------------
+
+  RISC OS file types are stored in bits 19..8 of the file load address.
+
+  To enable non-RISC OS systems to be used to store files without losing
+  file type information, a file naming convention was devised (initially
+  for use with NFS) such that a hexadecimal suffix of the form ,xyz
+  denoted the file type: e.g. BasicFile,ffb is a BASIC (0xffb) file.  This
+  naming convention is now also used by RISC OS emulators such as RPCEmu.
+
+  Mounting an ADFS disc with option ftsuffix=1 will cause appropriate file
+  type suffixes to be appended to file names read from a directory.  If the
+  ftsuffix option is zero or omitted, no file type suffixes will be added.
diff --git a/Documentation/filesystems/affs.txt b/Documentation/filesystems/affs.txt
new file mode 100644
index 000000000..a8f1a58e3
--- /dev/null
+++ b/Documentation/filesystems/affs.txt
@@ -0,0 +1,226 @@
+Overview of Amiga Filesystems
+=============================
+
+Not all varieties of the Amiga filesystems are supported for reading and
+writing. The Amiga currently knows six different filesystems:
+
+DOS\0		The old or original filesystem, not really suited for
+		hard disks and normally not used on them, either.
+		Supported read/write.
+
+DOS\1		The original Fast File System. Supported read/write.
+
+DOS\2		The old "international" filesystem. International means that
+		a bug has been fixed so that accented ("international") letters
+		in file names are case-insensitive, as they ought to be.
+		Supported read/write.
+
+DOS\3		The "international" Fast File System.  Supported read/write.
+
+DOS\4		The original filesystem with directory cache. The directory
+		cache speeds up directory accesses on floppies considerably,
+		but slows down file creation/deletion. Doesn't make much
+		sense on hard disks. Supported read only.
+
+DOS\5		The Fast File System with directory cache. Supported read only.
+
+All of the above filesystems allow block sizes from 512 to 32K bytes.
+Supported block sizes are: 512, 1024, 2048 and 4096 bytes. Larger blocks
+speed up almost everything at the expense of wasted disk space. The speed
+gain above 4K seems not really worth the price, so you don't lose too
+much here, either.
+
+The muFS (multi user File System) equivalents of the above file systems
+are supported, too.
+
+Mount options for the AFFS
+==========================
+
+protect		If this option is set, the protection bits cannot be altered.
+
+setuid[=uid]	This sets the owner of all files and directories in the file
+		system to uid or the uid of the current user, respectively.
+
+setgid[=gid]	Same as above, but for gid.
+
+mode=mode	Sets the mode flags to the given (octal) value, regardless
+		of the original permissions. Directories will get an x
+		permission if the corresponding r bit is set.
+		This is useful since most of the plain AmigaOS files
+		will map to 600.
+
+nofilenametruncate
+		The file system will return an error when filename exceeds
+		standard maximum filename length (30 characters).
+
+reserved=num	Sets the number of reserved blocks at the start of the
+		partition to num. You should never need this option.
+		Default is 2.
+
+root=block	Sets the block number of the root block. This should never
+		be necessary.
+
+bs=blksize	Sets the blocksize to blksize. Valid block sizes are 512,
+		1024, 2048 and 4096. Like the root option, this should
+		never be necessary, as the affs can figure it out itself.
+
+quiet		The file system will not return an error for disallowed
+		mode changes.
+
+verbose		The volume name, file system type and block size will
+		be written to the syslog when the filesystem is mounted.
+
+mufs		The filesystem is really a muFS, also it doesn't
+		identify itself as one. This option is necessary if
+		the filesystem wasn't formatted as muFS, but is used
+		as one.
+
+prefix=path	Path will be prefixed to every absolute path name of
+		symbolic links on an AFFS partition. Default = "/".
+		(See below.)
+
+volume=name	When symbolic links with an absolute path are created
+		on an AFFS partition, name will be prepended as the
+		volume name. Default = "" (empty string).
+		(See below.)
+
+Handling of the Users/Groups and protection flags
+=================================================
+
+Amiga -> Linux:
+
+The Amiga protection flags RWEDRWEDHSPARWED are handled as follows:
+
+  - R maps to r for user, group and others. On directories, R implies x.
+
+  - W maps to w.
+
+  - E maps to x.
+
+  - D is ignored.
+
+  - H, S and P are always retained and ignored under Linux.
+
+  - A is cleared when a file is written to.
+
+User id and group id will be used unless set[gu]id are given as mount
+options. Since most of the Amiga file systems are single user systems
+they will be owned by root. The root directory (the mount point) of the
+Amiga filesystem will be owned by the user who actually mounts the
+filesystem (the root directory doesn't have uid/gid fields).
+
+Linux -> Amiga:
+
+The Linux rwxrwxrwx file mode is handled as follows:
+
+  - r permission will allow R for user, group and others.
+
+  - w permission will allow W for user, group and others.
+
+  - x permission of the user will allow E for plain files.
+
+  - D will be allowed for user, group and others.
+
+  - All other flags (suid, sgid, ...) are ignored and will
+    not be retained.
+    
+Newly created files and directories will get the user and group ID
+of the current user and a mode according to the umask.
+
+Symbolic links
+==============
+
+Although the Amiga and Linux file systems resemble each other, there
+are some, not always subtle, differences. One of them becomes apparent
+with symbolic links. While Linux has a file system with exactly one
+root directory, the Amiga has a separate root directory for each
+file system (for example, partition, floppy disk, ...). With the Amiga,
+these entities are called "volumes". They have symbolic names which
+can be used to access them. Thus, symbolic links can point to a
+different volume. AFFS turns the volume name into a directory name
+and prepends the prefix path (see prefix option) to it.
+
+Example:
+You mount all your Amiga partitions under /amiga/<volume> (where
+<volume> is the name of the volume), and you give the option
+"prefix=/amiga/" when mounting all your AFFS partitions. (They
+might be "User", "WB" and "Graphics", the mount points /amiga/User,
+/amiga/WB and /amiga/Graphics). A symbolic link referring to
+"User:sc/include/dos/dos.h" will be followed to
+"/amiga/User/sc/include/dos/dos.h".
+
+Examples
+========
+
+Command line:
+    mount  Archive/Amiga/Workbench3.1.adf /mnt -t affs -o loop,verbose
+    mount  /dev/sda3 /Amiga -t affs
+
+/etc/fstab entry:
+    /dev/sdb5	/amiga/Workbench    affs    noauto,user,exec,verbose 0 0
+
+IMPORTANT NOTE
+==============
+
+If you boot Windows 95 (don't know about 3.x, 98 and NT) while you
+have an Amiga harddisk connected to your PC, it will overwrite
+the bytes 0x00dc..0x00df of block 0 with garbage, thus invalidating
+the Rigid Disk Block. Sheer luck has it that this is an unused
+area of the RDB, so only the checksum doesn't match anymore.
+Linux will ignore this garbage and recognize the RDB anyway, but
+before you connect that drive to your Amiga again, you must
+restore or repair your RDB. So please do make a backup copy of it
+before booting Windows!
+
+If the damage is already done, the following should fix the RDB
+(where <disk> is the device name).
+DO AT YOUR OWN RISK:
+
+  dd if=/dev/<disk> of=rdb.tmp count=1
+  cp rdb.tmp rdb.fixed
+  dd if=/dev/zero of=rdb.fixed bs=1 seek=220 count=4
+  dd if=rdb.fixed of=/dev/<disk>
+
+Bugs, Restrictions, Caveats
+===========================
+
+Quite a few things may not work as advertised. Not everything is
+tested, though several hundred MB have been read and written using
+this fs. For a most up-to-date list of bugs please consult
+fs/affs/Changes.
+
+By default, filenames are truncated to 30 characters without warning.
+'nofilenametruncate' mount option can change that behavior.
+
+Case is ignored by the affs in filename matching, but Linux shells
+do care about the case. Example (with /wb being an affs mounted fs):
+    rm /wb/WRONGCASE
+will remove /mnt/wrongcase, but
+    rm /wb/WR*
+will not since the names are matched by the shell.
+
+The block allocation is designed for hard disk partitions. If more
+than 1 process writes to a (small) diskette, the blocks are allocated
+in an ugly way (but the real AFFS doesn't do much better). This
+is also true when space gets tight.
+
+You cannot execute programs on an OFS (Old File System), since the
+program files cannot be memory mapped due to the 488 byte blocks.
+For the same reason you cannot mount an image on such a filesystem
+via the loopback device.
+
+The bitmap valid flag in the root block may not be accurate when the
+system crashes while an affs partition is mounted. There's currently
+no way to fix a garbled filesystem without an Amiga (disk validator)
+or manually (who would do this?). Maybe later.
+
+If you mount affs partitions on system startup, you may want to tell
+fsck that the fs should not be checked (place a '0' in the sixth field
+of /etc/fstab).
+
+It's not possible to read floppy disks with a normal PC or workstation
+due to an incompatibility with the Amiga floppy controller.
+
+If you are interested in an Amiga Emulator for Linux, look at
+
+http://web.archive.org/web/*/http://www.freiburg.linux.de/~uae/
diff --git a/Documentation/filesystems/afs.txt b/Documentation/filesystems/afs.txt
new file mode 100644
index 000000000..8c6ea7b41
--- /dev/null
+++ b/Documentation/filesystems/afs.txt
@@ -0,0 +1,258 @@
+			     ====================
+			     kAFS: AFS FILESYSTEM
+			     ====================
+
+Contents:
+
+ - Overview.
+ - Usage.
+ - Mountpoints.
+ - Dynamic root.
+ - Proc filesystem.
+ - The cell database.
+ - Security.
+ - The @sys substitution.
+
+
+========
+OVERVIEW
+========
+
+This filesystem provides a fairly simple secure AFS filesystem driver. It is
+under development and does not yet provide the full feature set.  The features
+it does support include:
+
+ (*) Security (currently only AFS kaserver and KerberosIV tickets).
+
+ (*) File reading and writing.
+
+ (*) Automounting.
+
+ (*) Local caching (via fscache).
+
+It does not yet support the following AFS features:
+
+ (*) pioctl() system call.
+
+
+===========
+COMPILATION
+===========
+
+The filesystem should be enabled by turning on the kernel configuration
+options:
+
+	CONFIG_AF_RXRPC		- The RxRPC protocol transport
+	CONFIG_RXKAD		- The RxRPC Kerberos security handler
+	CONFIG_AFS		- The AFS filesystem
+
+Additionally, the following can be turned on to aid debugging:
+
+	CONFIG_AF_RXRPC_DEBUG	- Permit AF_RXRPC debugging to be enabled
+	CONFIG_AFS_DEBUG	- Permit AFS debugging to be enabled
+
+They permit the debugging messages to be turned on dynamically by manipulating
+the masks in the following files:
+
+	/sys/module/af_rxrpc/parameters/debug
+	/sys/module/kafs/parameters/debug
+
+
+=====
+USAGE
+=====
+
+When inserting the driver modules the root cell must be specified along with a
+list of volume location server IP addresses:
+
+	modprobe rxrpc
+	modprobe kafs rootcell=cambridge.redhat.com:172.16.18.73:172.16.18.91
+
+The first module is the AF_RXRPC network protocol driver.  This provides the
+RxRPC remote operation protocol and may also be accessed from userspace.  See:
+
+	Documentation/networking/rxrpc.txt
+
+The second module is the kerberos RxRPC security driver, and the third module
+is the actual filesystem driver for the AFS filesystem.
+
+Once the module has been loaded, more modules can be added by the following
+procedure:
+
+	echo add grand.central.org 18.9.48.14:128.2.203.61:130.237.48.87 >/proc/fs/afs/cells
+
+Where the parameters to the "add" command are the name of a cell and a list of
+volume location servers within that cell, with the latter separated by colons.
+
+Filesystems can be mounted anywhere by commands similar to the following:
+
+	mount -t afs "%cambridge.redhat.com:root.afs." /afs
+	mount -t afs "#cambridge.redhat.com:root.cell." /afs/cambridge
+	mount -t afs "#root.afs." /afs
+	mount -t afs "#root.cell." /afs/cambridge
+
+Where the initial character is either a hash or a percent symbol depending on
+whether you definitely want a R/W volume (percent) or whether you'd prefer a
+R/O volume, but are willing to use a R/W volume instead (hash).
+
+The name of the volume can be suffixes with ".backup" or ".readonly" to
+specify connection to only volumes of those types.
+
+The name of the cell is optional, and if not given during a mount, then the
+named volume will be looked up in the cell specified during modprobe.
+
+Additional cells can be added through /proc (see later section).
+
+
+===========
+MOUNTPOINTS
+===========
+
+AFS has a concept of mountpoints. In AFS terms, these are specially formatted
+symbolic links (of the same form as the "device name" passed to mount).  kAFS
+presents these to the user as directories that have a follow-link capability
+(ie: symbolic link semantics).  If anyone attempts to access them, they will
+automatically cause the target volume to be mounted (if possible) on that site.
+
+Automatically mounted filesystems will be automatically unmounted approximately
+twenty minutes after they were last used.  Alternatively they can be unmounted
+directly with the umount() system call.
+
+Manually unmounting an AFS volume will cause any idle submounts upon it to be
+culled first.  If all are culled, then the requested volume will also be
+unmounted, otherwise error EBUSY will be returned.
+
+This can be used by the administrator to attempt to unmount the whole AFS tree
+mounted on /afs in one go by doing:
+
+	umount /afs
+
+
+============
+DYNAMIC ROOT
+============
+
+A mount option is available to create a serverless mount that is only usable
+for dynamic lookup.  Creating such a mount can be done by, for example:
+
+	mount -t afs none /afs -o dyn
+
+This creates a mount that just has an empty directory at the root.  Attempting
+to look up a name in this directory will cause a mountpoint to be created that
+looks up a cell of the same name, for example:
+
+	ls /afs/grand.central.org/
+
+
+===============
+PROC FILESYSTEM
+===============
+
+The AFS modules creates a "/proc/fs/afs/" directory and populates it:
+
+  (*) A "cells" file that lists cells currently known to the afs module and
+      their usage counts:
+
+	[root@andromeda ~]# cat /proc/fs/afs/cells
+	USE NAME
+	  3 cambridge.redhat.com
+
+  (*) A directory per cell that contains files that list volume location
+      servers, volumes, and active servers known within that cell.
+
+	[root@andromeda ~]# cat /proc/fs/afs/cambridge.redhat.com/servers
+	USE ADDR            STATE
+	  4 172.16.18.91        0
+	[root@andromeda ~]# cat /proc/fs/afs/cambridge.redhat.com/vlservers
+	ADDRESS
+	172.16.18.91
+	[root@andromeda ~]# cat /proc/fs/afs/cambridge.redhat.com/volumes
+	USE STT VLID[0]  VLID[1]  VLID[2]  NAME
+	  1 Val 20000000 20000001 20000002 root.afs
+
+
+=================
+THE CELL DATABASE
+=================
+
+The filesystem maintains an internal database of all the cells it knows and the
+IP addresses of the volume location servers for those cells.  The cell to which
+the system belongs is added to the database when modprobe is performed by the
+"rootcell=" argument or, if compiled in, using a "kafs.rootcell=" argument on
+the kernel command line.
+
+Further cells can be added by commands similar to the following:
+
+	echo add CELLNAME VLADDR[:VLADDR][:VLADDR]... >/proc/fs/afs/cells
+	echo add grand.central.org 18.9.48.14:128.2.203.61:130.237.48.87 >/proc/fs/afs/cells
+
+No other cell database operations are available at this time.
+
+
+========
+SECURITY
+========
+
+Secure operations are initiated by acquiring a key using the klog program.  A
+very primitive klog program is available at:
+
+	http://people.redhat.com/~dhowells/rxrpc/klog.c
+
+This should be compiled by:
+
+	make klog LDLIBS="-lcrypto -lcrypt -lkrb4 -lkeyutils"
+
+And then run as:
+
+	./klog
+
+Assuming it's successful, this adds a key of type RxRPC, named for the service
+and cell, eg: "afs@<cellname>".  This can be viewed with the keyctl program or
+by cat'ing /proc/keys:
+
+	[root@andromeda ~]# keyctl show
+	Session Keyring
+	       -3 --alswrv      0     0  keyring: _ses.3268
+		2 --alswrv      0     0   \_ keyring: _uid.0
+	111416553 --als--v      0     0   \_ rxrpc: afs@CAMBRIDGE.REDHAT.COM
+
+Currently the username, realm, password and proposed ticket lifetime are
+compiled in to the program.
+
+It is not required to acquire a key before using AFS facilities, but if one is
+not acquired then all operations will be governed by the anonymous user parts
+of the ACLs.
+
+If a key is acquired, then all AFS operations, including mounts and automounts,
+made by a possessor of that key will be secured with that key.
+
+If a file is opened with a particular key and then the file descriptor is
+passed to a process that doesn't have that key (perhaps over an AF_UNIX
+socket), then the operations on the file will be made with key that was used to
+open the file.
+
+
+=====================
+THE @SYS SUBSTITUTION
+=====================
+
+The list of up to 16 @sys substitutions for the current network namespace can
+be configured by writing a list to /proc/fs/afs/sysname:
+
+	[root@andromeda ~]# echo foo amd64_linux_26 >/proc/fs/afs/sysname
+
+or cleared entirely by writing an empty list:
+
+	[root@andromeda ~]# echo >/proc/fs/afs/sysname
+
+The current list for current network namespace can be retrieved by:
+
+	[root@andromeda ~]# cat /proc/fs/afs/sysname
+	foo
+	amd64_linux_26
+
+When @sys is being substituted for, each element of the list is tried in the
+order given.
+
+By default, the list will contain one item that conforms to the pattern
+"<arch>_linux_26", amd64 being the name for x86_64.
diff --git a/Documentation/filesystems/autofs-mount-control.txt b/Documentation/filesystems/autofs-mount-control.txt
new file mode 100644
index 000000000..45edad693
--- /dev/null
+++ b/Documentation/filesystems/autofs-mount-control.txt
@@ -0,0 +1,406 @@
+
+Miscellaneous Device control operations for the autofs kernel module
+====================================================================
+
+The problem
+===========
+
+There is a problem with active restarts in autofs (that is to say
+restarting autofs when there are busy mounts).
+
+During normal operation autofs uses a file descriptor opened on the
+directory that is being managed in order to be able to issue control
+operations. Using a file descriptor gives ioctl operations access to
+autofs specific information stored in the super block. The operations
+are things such as setting an autofs mount catatonic, setting the
+expire timeout and requesting expire checks. As is explained below,
+certain types of autofs triggered mounts can end up covering an autofs
+mount itself which prevents us being able to use open(2) to obtain a
+file descriptor for these operations if we don't already have one open.
+
+Currently autofs uses "umount -l" (lazy umount) to clear active mounts
+at restart. While using lazy umount works for most cases, anything that
+needs to walk back up the mount tree to construct a path, such as
+getcwd(2) and the proc file system /proc/<pid>/cwd, no longer works
+because the point from which the path is constructed has been detached
+from the mount tree.
+
+The actual problem with autofs is that it can't reconnect to existing
+mounts. Immediately one thinks of just adding the ability to remount
+autofs file systems would solve it, but alas, that can't work. This is
+because autofs direct mounts and the implementation of "on demand mount
+and expire" of nested mount trees have the file system mounted directly
+on top of the mount trigger directory dentry.
+
+For example, there are two types of automount maps, direct (in the kernel
+module source you will see a third type called an offset, which is just
+a direct mount in disguise) and indirect.
+
+Here is a master map with direct and indirect map entries:
+
+/-      /etc/auto.direct
+/test   /etc/auto.indirect
+
+and the corresponding map files:
+
+/etc/auto.direct:
+
+/automount/dparse/g6  budgie:/autofs/export1
+/automount/dparse/g1  shark:/autofs/export1
+and so on.
+
+/etc/auto.indirect:
+
+g1    shark:/autofs/export1
+g6    budgie:/autofs/export1
+and so on.
+
+For the above indirect map an autofs file system is mounted on /test and
+mounts are triggered for each sub-directory key by the inode lookup
+operation. So we see a mount of shark:/autofs/export1 on /test/g1, for
+example.
+
+The way that direct mounts are handled is by making an autofs mount on
+each full path, such as /automount/dparse/g1, and using it as a mount
+trigger. So when we walk on the path we mount shark:/autofs/export1 "on
+top of this mount point". Since these are always directories we can
+use the follow_link inode operation to trigger the mount.
+
+But, each entry in direct and indirect maps can have offsets (making
+them multi-mount map entries).
+
+For example, an indirect mount map entry could also be:
+
+g1  \
+   /        shark:/autofs/export5/testing/test \
+   /s1      shark:/autofs/export/testing/test/s1 \
+   /s2      shark:/autofs/export5/testing/test/s2 \
+   /s1/ss1  shark:/autofs/export1 \
+   /s2/ss2  shark:/autofs/export2
+
+and a similarly a direct mount map entry could also be:
+
+/automount/dparse/g1 \
+    /       shark:/autofs/export5/testing/test \
+    /s1     shark:/autofs/export/testing/test/s1 \
+    /s2     shark:/autofs/export5/testing/test/s2 \
+    /s1/ss1 shark:/autofs/export2 \
+    /s2/ss2 shark:/autofs/export2
+
+One of the issues with version 4 of autofs was that, when mounting an
+entry with a large number of offsets, possibly with nesting, we needed
+to mount and umount all of the offsets as a single unit. Not really a
+problem, except for people with a large number of offsets in map entries.
+This mechanism is used for the well known "hosts" map and we have seen
+cases (in 2.4) where the available number of mounts are exhausted or
+where the number of privileged ports available is exhausted.
+
+In version 5 we mount only as we go down the tree of offsets and
+similarly for expiring them which resolves the above problem. There is
+somewhat more detail to the implementation but it isn't needed for the
+sake of the problem explanation. The one important detail is that these
+offsets are implemented using the same mechanism as the direct mounts
+above and so the mount points can be covered by a mount.
+
+The current autofs implementation uses an ioctl file descriptor opened
+on the mount point for control operations. The references held by the
+descriptor are accounted for in checks made to determine if a mount is
+in use and is also used to access autofs file system information held
+in the mount super block. So the use of a file handle needs to be
+retained.
+
+
+The Solution
+============
+
+To be able to restart autofs leaving existing direct, indirect and
+offset mounts in place we need to be able to obtain a file handle
+for these potentially covered autofs mount points. Rather than just
+implement an isolated operation it was decided to re-implement the
+existing ioctl interface and add new operations to provide this
+functionality.
+
+In addition, to be able to reconstruct a mount tree that has busy mounts,
+the uid and gid of the last user that triggered the mount needs to be
+available because these can be used as macro substitution variables in
+autofs maps. They are recorded at mount request time and an operation
+has been added to retrieve them.
+
+Since we're re-implementing the control interface, a couple of other
+problems with the existing interface have been addressed. First, when
+a mount or expire operation completes a status is returned to the
+kernel by either a "send ready" or a "send fail" operation. The
+"send fail" operation of the ioctl interface could only ever send
+ENOENT so the re-implementation allows user space to send an actual
+status. Another expensive operation in user space, for those using
+very large maps, is discovering if a mount is present. Usually this
+involves scanning /proc/mounts and since it needs to be done quite
+often it can introduce significant overhead when there are many entries
+in the mount table. An operation to lookup the mount status of a mount
+point dentry (covered or not) has also been added.
+
+Current kernel development policy recommends avoiding the use of the
+ioctl mechanism in favor of systems such as Netlink. An implementation
+using this system was attempted to evaluate its suitability and it was
+found to be inadequate, in this case. The Generic Netlink system was
+used for this as raw Netlink would lead to a significant increase in
+complexity. There's no question that the Generic Netlink system is an
+elegant solution for common case ioctl functions but it's not a complete
+replacement probably because its primary purpose in life is to be a
+message bus implementation rather than specifically an ioctl replacement.
+While it would be possible to work around this there is one concern
+that lead to the decision to not use it. This is that the autofs
+expire in the daemon has become far to complex because umount
+candidates are enumerated, almost for no other reason than to "count"
+the number of times to call the expire ioctl. This involves scanning
+the mount table which has proved to be a big overhead for users with
+large maps. The best way to improve this is try and get back to the
+way the expire was done long ago. That is, when an expire request is
+issued for a mount (file handle) we should continually call back to
+the daemon until we can't umount any more mounts, then return the
+appropriate status to the daemon. At the moment we just expire one
+mount at a time. A Generic Netlink implementation would exclude this
+possibility for future development due to the requirements of the
+message bus architecture.
+
+
+autofs Miscellaneous Device mount control interface
+====================================================
+
+The control interface is opening a device node, typically /dev/autofs.
+
+All the ioctls use a common structure to pass the needed parameter
+information and return operation results:
+
+struct autofs_dev_ioctl {
+	__u32 ver_major;
+	__u32 ver_minor;
+	__u32 size;             /* total size of data passed in
+				 * including this struct */
+	__s32 ioctlfd;          /* automount command fd */
+
+	/* Command parameters */
+	union {
+		struct args_protover		protover;
+		struct args_protosubver		protosubver;
+		struct args_openmount		openmount;
+		struct args_ready		ready;
+		struct args_fail		fail;
+		struct args_setpipefd		setpipefd;
+		struct args_timeout		timeout;
+		struct args_requester		requester;
+		struct args_expire		expire;
+		struct args_askumount		askumount;
+		struct args_ismountpoint	ismountpoint;
+	};
+
+	char path[0];
+};
+
+The ioctlfd field is a mount point file descriptor of an autofs mount
+point. It is returned by the open call and is used by all calls except
+the check for whether a given path is a mount point, where it may
+optionally be used to check a specific mount corresponding to a given
+mount point file descriptor, and when requesting the uid and gid of the
+last successful mount on a directory within the autofs file system.
+
+The union is used to communicate parameters and results of calls made
+as described below.
+
+The path field is used to pass a path where it is needed and the size field
+is used account for the increased structure length when translating the
+structure sent from user space.
+
+This structure can be initialized before setting specific fields by using
+the void function call init_autofs_dev_ioctl(struct autofs_dev_ioctl *).
+
+All of the ioctls perform a copy of this structure from user space to
+kernel space and return -EINVAL if the size parameter is smaller than
+the structure size itself, -ENOMEM if the kernel memory allocation fails
+or -EFAULT if the copy itself fails. Other checks include a version check
+of the compiled in user space version against the module version and a
+mismatch results in a -EINVAL return. If the size field is greater than
+the structure size then a path is assumed to be present and is checked to
+ensure it begins with a "/" and is NULL terminated, otherwise -EINVAL is
+returned. Following these checks, for all ioctl commands except
+AUTOFS_DEV_IOCTL_VERSION_CMD, AUTOFS_DEV_IOCTL_OPENMOUNT_CMD and
+AUTOFS_DEV_IOCTL_CLOSEMOUNT_CMD the ioctlfd is validated and if it is
+not a valid descriptor or doesn't correspond to an autofs mount point
+an error of -EBADF, -ENOTTY or -EINVAL (not an autofs descriptor) is
+returned.
+
+
+The ioctls
+==========
+
+An example of an implementation which uses this interface can be seen
+in autofs version 5.0.4 and later in file lib/dev-ioctl-lib.c of the
+distribution tar available for download from kernel.org in directory
+/pub/linux/daemons/autofs/v5.
+
+The device node ioctl operations implemented by this interface are:
+
+
+AUTOFS_DEV_IOCTL_VERSION
+------------------------
+
+Get the major and minor version of the autofs device ioctl kernel module
+implementation. It requires an initialized struct autofs_dev_ioctl as an
+input parameter and sets the version information in the passed in structure.
+It returns 0 on success or the error -EINVAL if a version mismatch is
+detected.
+
+
+AUTOFS_DEV_IOCTL_PROTOVER_CMD and AUTOFS_DEV_IOCTL_PROTOSUBVER_CMD
+------------------------------------------------------------------
+
+Get the major and minor version of the autofs protocol version understood
+by loaded module. This call requires an initialized struct autofs_dev_ioctl
+with the ioctlfd field set to a valid autofs mount point descriptor
+and sets the requested version number in version field of struct args_protover
+or sub_version field of struct args_protosubver. These commands return
+0 on success or one of the negative error codes if validation fails.
+
+
+AUTOFS_DEV_IOCTL_OPENMOUNT and AUTOFS_DEV_IOCTL_CLOSEMOUNT
+----------------------------------------------------------
+
+Obtain and release a file descriptor for an autofs managed mount point
+path. The open call requires an initialized struct autofs_dev_ioctl with
+the path field set and the size field adjusted appropriately as well
+as the devid field of struct args_openmount set to the device number of
+the autofs mount. The device number can be obtained from the mount options
+shown in /proc/mounts. The close call requires an initialized struct
+autofs_dev_ioct with the ioctlfd field set to the descriptor obtained
+from the open call. The release of the file descriptor can also be done
+with close(2) so any open descriptors will also be closed at process exit.
+The close call is included in the implemented operations largely for
+completeness and to provide for a consistent user space implementation.
+
+
+AUTOFS_DEV_IOCTL_READY_CMD and AUTOFS_DEV_IOCTL_FAIL_CMD
+--------------------------------------------------------
+
+Return mount and expire result status from user space to the kernel.
+Both of these calls require an initialized struct autofs_dev_ioctl
+with the ioctlfd field set to the descriptor obtained from the open
+call and the token field of struct args_ready or struct args_fail set
+to the wait queue token number, received by user space in the foregoing
+mount or expire request. The status field of struct args_fail is set to
+the errno of the operation. It is set to 0 on success.
+
+
+AUTOFS_DEV_IOCTL_SETPIPEFD_CMD
+------------------------------
+
+Set the pipe file descriptor used for kernel communication to the daemon.
+Normally this is set at mount time using an option but when reconnecting
+to a existing mount we need to use this to tell the autofs mount about
+the new kernel pipe descriptor. In order to protect mounts against
+incorrectly setting the pipe descriptor we also require that the autofs
+mount be catatonic (see next call).
+
+The call requires an initialized struct autofs_dev_ioctl with the
+ioctlfd field set to the descriptor obtained from the open call and
+the pipefd field of struct args_setpipefd set to descriptor of the pipe.
+On success the call also sets the process group id used to identify the
+controlling process (eg. the owning automount(8) daemon) to the process
+group of the caller.
+
+
+AUTOFS_DEV_IOCTL_CATATONIC_CMD
+------------------------------
+
+Make the autofs mount point catatonic. The autofs mount will no longer
+issue mount requests, the kernel communication pipe descriptor is released
+and any remaining waits in the queue released.
+
+The call requires an initialized struct autofs_dev_ioctl with the
+ioctlfd field set to the descriptor obtained from the open call.
+
+
+AUTOFS_DEV_IOCTL_TIMEOUT_CMD
+----------------------------
+
+Set the expire timeout for mounts within an autofs mount point.
+
+The call requires an initialized struct autofs_dev_ioctl with the
+ioctlfd field set to the descriptor obtained from the open call.
+
+
+AUTOFS_DEV_IOCTL_REQUESTER_CMD
+------------------------------
+
+Return the uid and gid of the last process to successfully trigger a the
+mount on the given path dentry.
+
+The call requires an initialized struct autofs_dev_ioctl with the path
+field set to the mount point in question and the size field adjusted
+appropriately. Upon return the uid field of struct args_requester contains
+the uid and gid field the gid.
+
+When reconstructing an autofs mount tree with active mounts we need to
+re-connect to mounts that may have used the original process uid and
+gid (or string variations of them) for mount lookups within the map entry.
+This call provides the ability to obtain this uid and gid so they may be
+used by user space for the mount map lookups.
+
+
+AUTOFS_DEV_IOCTL_EXPIRE_CMD
+---------------------------
+
+Issue an expire request to the kernel for an autofs mount. Typically
+this ioctl is called until no further expire candidates are found.
+
+The call requires an initialized struct autofs_dev_ioctl with the
+ioctlfd field set to the descriptor obtained from the open call. In
+addition an immediate expire, independent of the mount timeout, can be
+requested by setting the how field of struct args_expire to 1. If no
+expire candidates can be found the ioctl returns -1 with errno set to
+EAGAIN.
+
+This call causes the kernel module to check the mount corresponding
+to the given ioctlfd for mounts that can be expired, issues an expire
+request back to the daemon and waits for completion.
+
+AUTOFS_DEV_IOCTL_ASKUMOUNT_CMD
+------------------------------
+
+Checks if an autofs mount point is in use.
+
+The call requires an initialized struct autofs_dev_ioctl with the
+ioctlfd field set to the descriptor obtained from the open call and
+it returns the result in the may_umount field of struct args_askumount,
+1 for busy and 0 otherwise.
+
+
+AUTOFS_DEV_IOCTL_ISMOUNTPOINT_CMD
+---------------------------------
+
+Check if the given path is a mountpoint.
+
+The call requires an initialized struct autofs_dev_ioctl. There are two
+possible variations. Both use the path field set to the path of the mount
+point to check and the size field adjusted appropriately. One uses the
+ioctlfd field to identify a specific mount point to check while the other
+variation uses the path and optionally in.type field of struct args_ismountpoint
+set to an autofs mount type. The call returns 1 if this is a mount point
+and sets out.devid field to the device number of the mount and out.magic
+field to the relevant super block magic number (described below) or 0 if
+it isn't a mountpoint. In both cases the the device number (as returned
+by new_encode_dev()) is returned in out.devid field.
+
+If supplied with a file descriptor we're looking for a specific mount,
+not necessarily at the top of the mounted stack. In this case the path
+the descriptor corresponds to is considered a mountpoint if it is itself
+a mountpoint or contains a mount, such as a multi-mount without a root
+mount. In this case we return 1 if the descriptor corresponds to a mount
+point and and also returns the super magic of the covering mount if there
+is one or 0 if it isn't a mountpoint.
+
+If a path is supplied (and the ioctlfd field is set to -1) then the path
+is looked up and is checked to see if it is the root of a mount. If a
+type is also given we are looking for a particular autofs mount and if
+a match isn't found a fail is returned. If the the located path is the
+root of a mount 1 is returned along with the super magic of the mount
+or 0 otherwise.
diff --git a/Documentation/filesystems/autofs.txt b/Documentation/filesystems/autofs.txt
new file mode 100644
index 000000000..373ad2585
--- /dev/null
+++ b/Documentation/filesystems/autofs.txt
@@ -0,0 +1,529 @@
+<head>
+<style> p { max-width:50em} ol, ul {max-width: 40em}</style>
+</head>
+
+autofs - how it works
+=====================
+
+Purpose
+-------
+
+The goal of autofs is to provide on-demand mounting and race free
+automatic unmounting of various other filesystems.  This provides two
+key advantages:
+
+1. There is no need to delay boot until all filesystems that
+   might be needed are mounted.  Processes that try to access those
+   slow filesystems might be delayed but other processes can
+   continue freely.  This is particularly important for
+   network filesystems (e.g. NFS) or filesystems stored on
+   media with a media-changing robot.
+
+2. The names and locations of filesystems can be stored in
+   a remote database and can change at any time.  The content
+   in that data base at the time of access will be used to provide
+   a target for the access.  The interpretation of names in the
+   filesystem can even be programmatic rather than database-backed,
+   allowing wildcards for example, and can vary based on the user who
+   first accessed a name.
+
+Context
+-------
+
+The "autofs" filesystem module is only one part of an autofs system.
+There also needs to be a user-space program which looks up names
+and mounts filesystems.  This will often be the "automount" program,
+though other tools including "systemd" can make use of "autofs".
+This document describes only the kernel module and the interactions
+required with any user-space program.  Subsequent text refers to this
+as the "automount daemon" or simply "the daemon".
+
+"autofs" is a Linux kernel module with provides the "autofs"
+filesystem type.  Several "autofs" filesystems can be mounted and they
+can each be managed separately, or all managed by the same daemon.
+
+Content
+-------
+
+An autofs filesystem can contain 3 sorts of objects: directories,
+symbolic links and mount traps.  Mount traps are directories with
+extra properties as described in the next section.
+
+Objects can only be created by the automount daemon: symlinks are
+created with a regular `symlink` system call, while directories and
+mount traps are created with `mkdir`.  The determination of whether a
+directory should be a mount trap or not is quite _ad hoc_, largely for
+historical reasons, and is determined in part by the
+*direct*/*indirect*/*offset* mount options, and the *maxproto* mount option.
+
+If neither the *direct* or *offset* mount options are given (so the
+mount is considered to be *indirect*), then the root directory is
+always a regular directory, otherwise it is a mount trap when it is
+empty and a regular directory when not empty.  Note that *direct* and
+*offset* are treated identically so a concise summary is that the root
+directory is a mount trap only if the filesystem is mounted *direct*
+and the root is empty.
+
+Directories created in the root directory are mount traps only if the
+filesystem is mounted *indirect* and they are empty.
+
+Directories further down the tree depend on the *maxproto* mount
+option and particularly whether it is less than five or not.
+When *maxproto* is five, no directories further down the
+tree are ever mount traps, they are always regular directories.  When
+the *maxproto* is four (or three), these directories are mount traps
+precisely when they are empty.
+
+So: non-empty (i.e. non-leaf) directories are never mount traps. Empty
+directories are sometimes mount traps, and sometimes not depending on
+where in the tree they are (root, top level, or lower), the *maxproto*,
+and whether the mount was *indirect* or not.
+
+Mount Traps
+---------------
+
+A core element of the implementation of autofs is the Mount Traps
+which are provided by the Linux VFS.  Any directory provided by a
+filesystem can be designated as a trap.  This involves two separate
+features that work together to allow autofs to do its job.
+
+**DCACHE_NEED_AUTOMOUNT**
+
+If a dentry has the DCACHE_NEED_AUTOMOUNT flag set (which gets set if
+the inode has S_AUTOMOUNT set, or can be set directly) then it is
+(potentially) a mount trap.  Any access to this directory beyond a
+"`stat`" will (normally) cause the `d_op->d_automount()` dentry operation
+to be called. The task of this method is to find the filesystem that
+should be mounted on the directory and to return it.  The VFS is
+responsible for actually mounting the root of this filesystem on the
+directory.
+
+autofs doesn't find the filesystem itself but sends a message to the
+automount daemon asking it to find and mount the filesystem.  The
+autofs `d_automount` method then waits for the daemon to report that
+everything is ready.  It will then return "`NULL`" indicating that the
+mount has already happened.  The VFS doesn't try to mount anything but
+follows down the mount that is already there.
+
+This functionality is sufficient for some users of mount traps such
+as NFS which creates traps so that mountpoints on the server can be
+reflected on the client.  However it is not sufficient for autofs.  As
+mounting onto a directory is considered to be "beyond a `stat`", the
+automount daemon would not be able to mount a filesystem on the 'trap'
+directory without some way to avoid getting caught in the trap.  For
+that purpose there is another flag.
+
+**DCACHE_MANAGE_TRANSIT**
+
+If a dentry has DCACHE_MANAGE_TRANSIT set then two very different but
+related behaviors are invoked, both using the `d_op->d_manage()`
+dentry operation.
+
+Firstly, before checking to see if any filesystem is mounted on the
+directory, d_manage() will be called with the `rcu_walk` parameter set
+to `false`.  It may return one of three things:
+
+-  A return value of zero indicates that there is nothing special
+   about this dentry and normal checks for mounts and automounts
+   should proceed.
+
+   autofs normally returns zero, but first waits for any
+   expiry (automatic unmounting of the mounted filesystem) to
+   complete.  This avoids races.
+
+-  A return value of `-EISDIR` tells the VFS to ignore any mounts
+   on the directory and to not consider calling `->d_automount()`.
+   This effectively disables the **DCACHE_NEED_AUTOMOUNT** flag
+   causing the directory not be a mount trap after all.
+
+   autofs returns this if it detects that the process performing the
+   lookup is the automount daemon and that the mount has been
+   requested but has not yet completed.  How it determines this is
+   discussed later.  This allows the automount daemon not to get
+   caught in the mount trap.
+
+   There is a subtlety here.  It is possible that a second autofs
+   filesystem can be mounted below the first and for both of them to
+   be managed by the same daemon.  For the daemon to be able to mount
+   something on the second it must be able to "walk" down past the
+   first.  This means that d_manage cannot *always* return -EISDIR for
+   the automount daemon.  It must only return it when a mount has
+   been requested, but has not yet completed.
+
+   `d_manage` also returns `-EISDIR` if the dentry shouldn't be a
+   mount trap, either because it is a symbolic link or because it is
+   not empty.
+
+-  Any other negative value is treated as an error and returned
+   to the caller.
+
+   autofs can return
+
+   - -ENOENT if the automount daemon failed to mount anything,
+   - -ENOMEM if it ran out of memory,
+   - -EINTR if a signal arrived while waiting for expiry to
+     complete
+   - or any other error sent down by the automount daemon.
+
+
+The second use case only occurs during an "RCU-walk" and so `rcu_walk`
+will be set.
+
+An RCU-walk is a fast and lightweight process for walking down a
+filename path (i.e. it is like running on tip-toes).  RCU-walk cannot
+cope with all situations so when it finds a difficulty it falls back
+to "REF-walk", which is slower but more robust.
+
+RCU-walk will never call `->d_automount`; the filesystems must already
+be mounted or RCU-walk cannot handle the path.
+To determine if a mount-trap is safe for RCU-walk mode it calls
+`->d_manage()` with `rcu_walk` set to `true`.
+
+In this case `d_manage()` must avoid blocking and should avoid taking
+spinlocks if at all possible.  Its sole purpose is to determine if it
+would be safe to follow down into any mounted directory and the only
+reason that it might not be is if an expiry of the mount is
+underway.
+
+In the `rcu_walk` case, `d_manage()` cannot return -EISDIR to tell the
+VFS that this is a directory that doesn't require d_automount.  If
+`rcu_walk` sees a dentry with DCACHE_NEED_AUTOMOUNT set but nothing
+mounted, it *will* fall back to REF-walk.  `d_manage()` cannot make the
+VFS remain in RCU-walk mode, but can only tell it to get out of
+RCU-walk mode by returning `-ECHILD`.
+
+So `d_manage()`, when called with `rcu_walk` set, should either return
+-ECHILD if there is any reason to believe it is unsafe to end the
+mounted filesystem, and otherwise should return 0.
+
+autofs will return `-ECHILD` if an expiry of the filesystem has been
+initiated or is being considered, otherwise it returns 0.
+
+
+Mountpoint expiry
+-----------------
+
+The VFS has a mechanism for automatically expiring unused mounts,
+much as it can expire any unused dentry information from the dcache.
+This is guided by the MNT_SHRINKABLE flag.  This only applies to
+mounts that were created by `d_automount()` returning a filesystem to be
+mounted.  As autofs doesn't return such a filesystem but leaves the
+mounting to the automount daemon, it must involve the automount daemon
+in unmounting as well.  This also means that autofs has more control
+of expiry.
+
+The VFS also supports "expiry" of mounts using the MNT_EXPIRE flag to
+the `umount` system call.  Unmounting with MNT_EXPIRE will fail unless
+a previous attempt had been made, and the filesystem has been inactive
+and untouched since that previous attempt.  autofs does not depend on
+this but has its own internal tracking of whether filesystems were
+recently used.  This allows individual names in the autofs directory
+to expire separately.
+
+With version 4 of the protocol, the automount daemon can try to
+unmount any filesystems mounted on the autofs filesystem or remove any
+symbolic links or empty directories any time it likes.  If the unmount
+or removal is successful the filesystem will be returned to the state
+it was before the mount or creation, so that any access of the name
+will trigger normal auto-mount processing.  In particlar, `rmdir` and
+`unlink` do not leave negative entries in the dcache as a normal
+filesystem would, so an attempt to access a recently-removed object is
+passed to autofs for handling.
+
+With version 5, this is not safe except for unmounting from top-level
+directories.  As lower-level directories are never mount traps, other
+processes will see an empty directory as soon as the filesystem is
+unmounted.  So it is generally safest to use the autofs expiry
+protocol described below.
+
+Normally the daemon only wants to remove entries which haven't been
+used for a while.  For this purpose autofs maintains a "`last_used`"
+time stamp on each directory or symlink.  For symlinks it genuinely
+does record the last time the symlink was "used" or followed to find
+out where it points to.  For directories the field is a slight
+misnomer.  It actually records the last time that autofs checked if
+the directory or one of its descendents was busy and found that it
+was.  This is just as useful and doesn't require updating the field so
+often.
+
+The daemon is able to ask autofs if anything is due to be expired,
+using an `ioctl` as discussed later.  For a *direct* mount, autofs
+considers if the entire mount-tree can be unmounted or not.  For an
+*indirect* mount, autofs considers each of the names in the top level
+directory to determine if any of those can be unmounted and cleaned
+up.
+
+There is an option with indirect mounts to consider each of the leaves
+that has been mounted on instead of considering the top-level names.
+This is intended for compatability with version 4 of autofs and should
+be considered as deprecated.
+
+When autofs considers a directory it checks the `last_used` time and
+compares it with the "timeout" value set when the filesystem was
+mounted, though this check is ignored in some cases. It also checks if
+the directory or anything below it is in use.  For symbolic links,
+only the `last_used` time is ever considered.
+
+If both appear to support expiring the directory or symlink, an action
+is taken.
+
+There are two ways to ask autofs to consider expiry.  The first is to
+use the **AUTOFS_IOC_EXPIRE** ioctl.  This only works for indirect
+mounts.  If it finds something in the root directory to expire it will
+return the name of that thing.  Once a name has been returned the
+automount daemon needs to unmount any filesystems mounted below the
+name normally.  As described above, this is unsafe for non-toplevel
+mounts in a version-5 autofs.  For this reason the current `automountd`
+does not use this ioctl.
+
+The second mechanism uses either the **AUTOFS_DEV_IOCTL_EXPIRE_CMD** or
+the **AUTOFS_IOC_EXPIRE_MULTI** ioctl.  This will work for both direct and
+indirect mounts.  If it selects an object to expire, it will notify
+the daemon using the notification mechanism described below.  This
+will block until the daemon acknowledges the expiry notification.
+This implies that the "`EXPIRE`" ioctl must be sent from a different
+thread than the one which handles notification.
+
+While the ioctl is blocking, the entry is marked as "expiring" and
+`d_manage` will block until the daemon affirms that the unmount has
+completed (together with removing any directories that might have been
+necessary), or has been aborted.
+
+Communicating with autofs: detecting the daemon
+-----------------------------------------------
+
+There are several forms of communication between the automount daemon
+and the filesystem.  As we have already seen, the daemon can create and
+remove directories and symlinks using normal filesystem operations.
+autofs knows whether a process requesting some operation is the daemon
+or not based on its process-group id number (see getpgid(1)).
+
+When an autofs filesystem is mounted the pgid of the mounting
+processes is recorded unless the "pgrp=" option is given, in which
+case that number is recorded instead.  Any request arriving from a
+process in that process group is considered to come from the daemon.
+If the daemon ever has to be stopped and restarted a new pgid can be
+provided through an ioctl as will be described below.
+
+Communicating with autofs: the event pipe
+-----------------------------------------
+
+When an autofs filesystem is mounted, the 'write' end of a pipe must
+be passed using the 'fd=' mount option.  autofs will write
+notification messages to this pipe for the daemon to respond to.
+For version 5, the format of the message is:
+
+        struct autofs_v5_packet {
+                int proto_version;                /* Protocol version */
+                int type;                        /* Type of packet */
+                autofs_wqt_t wait_queue_token;
+                __u32 dev;
+                __u64 ino;
+                __u32 uid;
+                __u32 gid;
+                __u32 pid;
+                __u32 tgid;
+                __u32 len;
+                char name[NAME_MAX+1];
+        };
+
+where the type is one of
+
+        autofs_ptype_missing_indirect
+        autofs_ptype_expire_indirect
+        autofs_ptype_missing_direct
+        autofs_ptype_expire_direct
+
+so messages can indicate that a name is missing (something tried to
+access it but it isn't there) or that it has been selected for expiry.
+
+The pipe will be set to "packet mode" (equivalent to passing
+`O_DIRECT`) to _pipe2(2)_ so that a read from the pipe will return at
+most one packet, and any unread portion of a packet will be discarded.
+
+The `wait_queue_token` is a unique number which can identify a
+particular request to be acknowledged.  When a message is sent over
+the pipe the affected dentry is marked as either "active" or
+"expiring" and other accesses to it block until the message is
+acknowledged using one of the ioctls below and the relevant
+`wait_queue_token`.
+
+Communicating with autofs: root directory ioctls
+------------------------------------------------
+
+The root directory of an autofs filesystem will respond to a number of
+ioctls.  The process issuing the ioctl must have the CAP_SYS_ADMIN
+capability, or must be the automount daemon.
+
+The available ioctl commands are:
+
+- **AUTOFS_IOC_READY**: a notification has been handled.  The argument
+    to the ioctl command is the "wait_queue_token" number
+    corresponding to the notification being acknowledged.
+- **AUTOFS_IOC_FAIL**: similar to above, but indicates failure with
+    the error code `ENOENT`.
+- **AUTOFS_IOC_CATATONIC**: Causes the autofs to enter "catatonic"
+    mode meaning that it stops sending notifications to the daemon.
+    This mode is also entered if a write to the pipe fails.
+- **AUTOFS_IOC_PROTOVER**:  This returns the protocol version in use.
+- **AUTOFS_IOC_PROTOSUBVER**: Returns the protocol sub-version which
+    is really a version number for the implementation.  It is
+    currently 2.
+- **AUTOFS_IOC_SETTIMEOUT**:  This passes a pointer to an unsigned
+    long.  The value is used to set the timeout for expiry, and
+    the current timeout value is stored back through the pointer.
+- **AUTOFS_IOC_ASKUMOUNT**:  Returns, in the pointed-to `int`, 1 if
+    the filesystem could be unmounted.  This is only a hint as
+    the situation could change at any instant.  This call can be
+    use to avoid a more expensive full unmount attempt.
+- **AUTOFS_IOC_EXPIRE**: as described above, this asks if there is
+    anything suitable to expire.  A pointer to a packet:
+
+        struct autofs_packet_expire_multi {
+                int proto_version;              /* Protocol version */
+                int type;                       /* Type of packet */
+                autofs_wqt_t wait_queue_token;
+                int len;
+                char name[NAME_MAX+1];
+        };
+
+     is required.  This is filled in with the name of something
+     that can be unmounted or removed.  If nothing can be expired,
+     `errno` is set to `EAGAIN`.  Even though a `wait_queue_token`
+     is present in the structure, no "wait queue" is established
+     and no acknowledgment is needed.
+- **AUTOFS_IOC_EXPIRE_MULTI**:  This is similar to
+     **AUTOFS_IOC_EXPIRE** except that it causes notification to be
+     sent to the daemon, and it blocks until the daemon acknowledges.
+     The argument is an integer which can contain two different flags.
+
+     **AUTOFS_EXP_IMMEDIATE** causes `last_used` time to be ignored
+     and objects are expired if the are not in use.
+
+     **AUTOFS_EXP_LEAVES** will select a leaf rather than a top-level
+     name to expire.  This is only safe when *maxproto* is 4.
+
+Communicating with autofs: char-device ioctls
+---------------------------------------------
+
+It is not always possible to open the root of an autofs filesystem,
+particularly a *direct* mounted filesystem.  If the automount daemon
+is restarted there is no way for it to regain control of existing
+mounts using any of the above communication channels.  To address this
+need there is a "miscellaneous" character device (major 10, minor 235)
+which can be used to communicate directly with the autofs filesystem.
+It requires CAP_SYS_ADMIN for access.
+
+The `ioctl`s that can be used on this device are described in a separate
+document `autofs-mount-control.txt`, and are summarized briefly here.
+Each ioctl is passed a pointer to an `autofs_dev_ioctl` structure:
+
+        struct autofs_dev_ioctl {
+                __u32 ver_major;
+                __u32 ver_minor;
+                __u32 size;             /* total size of data passed in
+                                         * including this struct */
+                __s32 ioctlfd;          /* automount command fd */
+
+		/* Command parameters */
+		union {
+			struct args_protover		protover;
+			struct args_protosubver		protosubver;
+			struct args_openmount		openmount;
+			struct args_ready		ready;
+			struct args_fail		fail;
+			struct args_setpipefd		setpipefd;
+			struct args_timeout		timeout;
+			struct args_requester		requester;
+			struct args_expire		expire;
+			struct args_askumount		askumount;
+			struct args_ismountpoint	ismountpoint;
+		};
+
+                char path[0];
+        };
+
+For the **OPEN_MOUNT** and **IS_MOUNTPOINT** commands, the target
+filesystem is identified by the `path`.  All other commands identify
+the filesystem by the `ioctlfd` which is a file descriptor open on the
+root, and which can be returned by **OPEN_MOUNT**.
+
+The `ver_major` and `ver_minor` are in/out parameters which check that
+the requested version is supported, and report the maximum version
+that the kernel module can support.
+
+Commands are:
+
+- **AUTOFS_DEV_IOCTL_VERSION_CMD**: does nothing, except validate and
+    set version numbers.
+- **AUTOFS_DEV_IOCTL_OPENMOUNT_CMD**: return an open file descriptor
+    on the root of an autofs filesystem.  The filesystem is identified
+    by name and device number, which is stored in `openmount.devid`.
+    Device numbers for existing filesystems can be found in
+    `/proc/self/mountinfo`.
+- **AUTOFS_DEV_IOCTL_CLOSEMOUNT_CMD**: same as `close(ioctlfd)`.
+- **AUTOFS_DEV_IOCTL_SETPIPEFD_CMD**: if the filesystem is in
+    catatonic mode, this can provide the write end of a new pipe
+    in `setpipefd.pipefd` to re-establish communication with a daemon.
+    The process group of the calling process is used to identify the
+    daemon.
+- **AUTOFS_DEV_IOCTL_REQUESTER_CMD**: `path` should be a
+    name within the filesystem that has been auto-mounted on.
+    On successful return, `requester.uid` and `requester.gid` will be
+    the UID and GID of the process which triggered that mount.
+- **AUTOFS_DEV_IOCTL_ISMOUNTPOINT_CMD**: Check if path is a
+    mountpoint of a particular type - see separate documentation for
+    details.
+- **AUTOFS_DEV_IOCTL_PROTOVER_CMD**:
+- **AUTOFS_DEV_IOCTL_PROTOSUBVER_CMD**:
+- **AUTOFS_DEV_IOCTL_READY_CMD**:
+- **AUTOFS_DEV_IOCTL_FAIL_CMD**:
+- **AUTOFS_DEV_IOCTL_CATATONIC_CMD**:
+- **AUTOFS_DEV_IOCTL_TIMEOUT_CMD**:
+- **AUTOFS_DEV_IOCTL_EXPIRE_CMD**:
+- **AUTOFS_DEV_IOCTL_ASKUMOUNT_CMD**:  These all have the same
+    function as the similarly named **AUTOFS_IOC** ioctls, except
+    that **FAIL** can be given an explicit error number in `fail.status`
+    instead of assuming `ENOENT`, and this **EXPIRE** command
+    corresponds to **AUTOFS_IOC_EXPIRE_MULTI**.
+
+Catatonic mode
+--------------
+
+As mentioned, an autofs mount can enter "catatonic" mode.  This
+happens if a write to the notification pipe fails, or if it is
+explicitly requested by an `ioctl`.
+
+When entering catatonic mode, the pipe is closed and any pending
+notifications are acknowledged with the error `ENOENT`.
+
+Once in catatonic mode attempts to access non-existing names will
+result in `ENOENT` while attempts to access existing directories will
+be treated in the same way as if they came from the daemon, so mount
+traps will not fire.
+
+When the filesystem is mounted a _uid_ and _gid_ can be given which
+set the ownership of directories and symbolic links.  When the
+filesystem is in catatonic mode, any process with a matching UID can
+create directories or symlinks in the root directory, but not in other
+directories.
+
+Catatonic mode can only be left via the
+**AUTOFS_DEV_IOCTL_OPENMOUNT_CMD** ioctl on the `/dev/autofs`.
+
+autofs, name spaces, and shared mounts
+--------------------------------------
+
+With bind mounts and name spaces it is possible for an autofs
+filesystem to appear at multiple places in one or more filesystem
+name spaces.  For this to work sensibly, the autofs filesystem should
+always be mounted "shared". e.g.
+
+> `mount --make-shared /autofs/mount/point`
+
+The automount daemon is only able to manage a single mount location for
+an autofs filesystem and if mounts on that are not 'shared', other
+locations will not behave as expected.  In particular access to those
+other locations will likely result in the `ELOOP` error
+
+> Too many levels of symbolic links
diff --git a/Documentation/filesystems/automount-support.txt b/Documentation/filesystems/automount-support.txt
new file mode 100644
index 000000000..b0afd3d55
--- /dev/null
+++ b/Documentation/filesystems/automount-support.txt
@@ -0,0 +1,93 @@
+Support is available for filesystems that wish to do automounting
+support (such as kAFS which can be found in fs/afs/ and NFS in
+fs/nfs/). This facility includes allowing in-kernel mounts to be
+performed and mountpoint degradation to be requested. The latter can
+also be requested by userspace.
+
+
+======================
+IN-KERNEL AUTOMOUNTING
+======================
+
+See section "Mount Traps" of  Documentation/filesystems/autofs.txt
+
+Then from userspace, you can just do something like:
+
+	[root@andromeda root]# mount -t afs \#root.afs. /afs
+	[root@andromeda root]# ls /afs
+	asd  cambridge  cambridge.redhat.com  grand.central.org
+	[root@andromeda root]# ls /afs/cambridge
+	afsdoc
+	[root@andromeda root]# ls /afs/cambridge/afsdoc/
+	ChangeLog  html  LICENSE  pdf  RELNOTES-1.2.2
+
+And then if you look in the mountpoint catalogue, you'll see something like:
+
+	[root@andromeda root]# cat /proc/mounts
+	...
+	#root.afs. /afs afs rw 0 0
+	#root.cell. /afs/cambridge.redhat.com afs rw 0 0
+	#afsdoc. /afs/cambridge.redhat.com/afsdoc afs rw 0 0
+
+
+===========================
+AUTOMATIC MOUNTPOINT EXPIRY
+===========================
+
+Automatic expiration of mountpoints is easy, provided you've mounted the
+mountpoint to be expired in the automounting procedure outlined separately.
+
+To do expiration, you need to follow these steps:
+
+ (1) Create at least one list off which the vfsmounts to be expired can be
+     hung.
+
+ (2) When a new mountpoint is created in the ->d_automount method, add
+     the mnt to the list using mnt_set_expiry()
+             mnt_set_expiry(newmnt, &afs_vfsmounts);
+
+ (3) When you want mountpoints to be expired, call mark_mounts_for_expiry()
+     with a pointer to this list. This will process the list, marking every
+     vfsmount thereon for potential expiry on the next call.
+
+     If a vfsmount was already flagged for expiry, and if its usage count is 1
+     (it's only referenced by its parent vfsmount), then it will be deleted
+     from the namespace and thrown away (effectively unmounted).
+
+     It may prove simplest to simply call this at regular intervals, using
+     some sort of timed event to drive it.
+
+The expiration flag is cleared by calls to mntput. This means that expiration
+will only happen on the second expiration request after the last time the
+mountpoint was accessed.
+
+If a mountpoint is moved, it gets removed from the expiration list. If a bind
+mount is made on an expirable mount, the new vfsmount will not be on the
+expiration list and will not expire.
+
+If a namespace is copied, all mountpoints contained therein will be copied,
+and the copies of those that are on an expiration list will be added to the
+same expiration list.
+
+
+=======================
+USERSPACE DRIVEN EXPIRY
+=======================
+
+As an alternative, it is possible for userspace to request expiry of any
+mountpoint (though some will be rejected - the current process's idea of the
+rootfs for example). It does this by passing the MNT_EXPIRE flag to
+umount(). This flag is considered incompatible with MNT_FORCE and MNT_DETACH.
+
+If the mountpoint in question is in referenced by something other than
+umount() or its parent mountpoint, an EBUSY error will be returned and the
+mountpoint will not be marked for expiration or unmounted.
+
+If the mountpoint was not already marked for expiry at that time, an EAGAIN
+error will be given and it won't be unmounted.
+
+Otherwise if it was already marked and it wasn't referenced, unmounting will
+take place as usual.
+
+Again, the expiration flag is cleared every time anything other than umount()
+looks at a mountpoint.
diff --git a/Documentation/filesystems/befs.txt b/Documentation/filesystems/befs.txt
new file mode 100644
index 000000000..da45e6c84
--- /dev/null
+++ b/Documentation/filesystems/befs.txt
@@ -0,0 +1,117 @@
+BeOS filesystem for Linux
+
+Document last updated: Dec 6, 2001
+
+WARNING
+=======
+Make sure you understand that this is alpha software.  This means that the
+implementation is neither complete nor well-tested. 
+
+I DISCLAIM ALL RESPONSIBILITY FOR ANY POSSIBLE BAD EFFECTS OF THIS CODE!
+
+LICENSE
+=====
+This software is covered by the GNU General Public License. 
+See the file COPYING for the complete text of the license.
+Or the GNU website: <http://www.gnu.org/licenses/licenses.html>
+
+AUTHOR
+=====
+The largest part of the code written by Will Dyson <will_dyson@pobox.com>
+He has been working on the code since Aug 13, 2001. See the changelog for
+details.
+
+Original Author: Makoto Kato <m_kato@ga2.so-net.ne.jp>
+His original code can still be found at:
+<http://hp.vector.co.jp/authors/VA008030/bfs/>
+Does anyone know of a more current email address for Makoto? He doesn't
+respond to the address given above...
+
+This filesystem doesn't have a maintainer.
+
+WHAT IS THIS DRIVER?
+==================
+This module implements the native filesystem of BeOS http://www.beincorporated.com/ 
+for the linux 2.4.1 and later kernels. Currently it is a read-only
+implementation.
+
+Which is it, BFS or BEFS?
+================
+Be, Inc said, "BeOS Filesystem is officially called BFS, not BeFS". 
+But Unixware Boot Filesystem is called bfs, too. And they are already in
+the kernel. Because of this naming conflict, on Linux the BeOS
+filesystem is called befs.
+
+HOW TO INSTALL
+==============
+step 1.  Install the BeFS  patch into the source code tree of linux.
+
+Apply the patchfile to your kernel source tree.
+Assuming that your kernel source is in /foo/bar/linux and the patchfile
+is called patch-befs-xxx, you would do the following:
+
+	cd /foo/bar/linux
+	patch -p1 < /path/to/patch-befs-xxx
+
+if the patching step fails (i.e. there are rejected hunks), you can try to
+figure it out yourself (it shouldn't be hard), or mail the maintainer 
+(Will Dyson <will_dyson@pobox.com>) for help.
+
+step 2.  Configuration & make kernel
+
+The linux kernel has many compile-time options. Most of them are beyond the
+scope of this document. I suggest the Kernel-HOWTO document as a good general
+reference on this topic. http://www.linuxdocs.org/HOWTOs/Kernel-HOWTO-4.html 
+
+However, to use the BeFS module, you must enable it at configure time.
+
+	cd /foo/bar/linux
+	make menuconfig (or xconfig)
+
+The BeFS module is not a standard part of the linux kernel, so you must first
+enable support for experimental code under the "Code maturity level" menu.
+
+Then, under the "Filesystems" menu will be an option called "BeFS
+filesystem (experimental)", or something like that. Enable that option
+(it is fine to make it a module).
+
+Save your kernel configuration and then build your kernel.
+
+step 3.  Install
+
+See the kernel howto <http://www.linux.com/howto/Kernel-HOWTO.html> for
+instructions on this critical step.
+
+USING BFS
+=========
+To use the BeOS filesystem, use filesystem type 'befs'.
+
+ex)
+    mount -t befs /dev/fd0 /beos
+
+MOUNT OPTIONS
+=============
+uid=nnn        All files in the partition will be owned by user id nnn.
+gid=nnn	       All files in the partition will be in group nnn.
+iocharset=xxx  Use xxx as the name of the NLS translation table.
+debug          The driver will output debugging information to the syslog.
+
+HOW TO GET LASTEST VERSION
+==========================
+
+The latest version is currently available at:
+<http://befs-driver.sourceforge.net/>
+
+ANY KNOWN BUGS?
+===========
+As of Jan 20, 2002:
+	
+	None
+
+SPECIAL THANKS
+==============
+Dominic Giampalo ... Writing "Practical file system design with Be filesystem"
+Hiroyuki Yamada  ... Testing LinuxPPC.
+
+
+
diff --git a/Documentation/filesystems/bfs.txt b/Documentation/filesystems/bfs.txt
new file mode 100644
index 000000000..843ce91a2
--- /dev/null
+++ b/Documentation/filesystems/bfs.txt
@@ -0,0 +1,57 @@
+BFS FILESYSTEM FOR LINUX
+========================
+
+The BFS filesystem is used by SCO UnixWare OS for the /stand slice, which
+usually contains the kernel image and a few other files required for the
+boot process.
+
+In order to access /stand partition under Linux you obviously need to
+know the partition number and the kernel must support UnixWare disk slices
+(CONFIG_UNIXWARE_DISKLABEL config option). However BFS support does not
+depend on having UnixWare disklabel support because one can also mount
+BFS filesystem via loopback:
+
+# losetup /dev/loop0 stand.img
+# mount -t bfs /dev/loop0 /mnt/stand
+
+where stand.img is a file containing the image of BFS filesystem. 
+When you have finished using it and umounted you need to also deallocate
+/dev/loop0 device by:
+
+# losetup -d /dev/loop0
+
+You can simplify mounting by just typing:
+
+# mount -t bfs -o loop stand.img /mnt/stand
+
+this will allocate the first available loopback device (and load loop.o 
+kernel module if necessary) automatically. If the loopback driver is not
+loaded automatically, make sure that you have compiled the module and
+that modprobe is functioning. Beware that umount will not deallocate
+/dev/loopN device if /etc/mtab file on your system is a symbolic link to
+/proc/mounts. You will need to do it manually using "-d" switch of
+losetup(8). Read losetup(8) manpage for more info.
+
+To create the BFS image under UnixWare you need to find out first which
+slice contains it. The command prtvtoc(1M) is your friend:
+
+# prtvtoc /dev/rdsk/c0b0t0d0s0
+
+(assuming your root disk is on target=0, lun=0, bus=0, controller=0). Then you
+look for the slice with tag "STAND", which is usually slice 10. With this
+information you can use dd(1) to create the BFS image:
+
+# umount /stand
+# dd if=/dev/rdsk/c0b0t0d0sa of=stand.img bs=512
+
+Just in case, you can verify that you have done the right thing by checking
+the magic number:
+
+# od -Ad -tx4 stand.img | more
+
+The first 4 bytes should be 0x1badface.
+
+If you have any patches, questions or suggestions regarding this BFS
+implementation please contact the author:
+
+Tigran Aivazian <aivazian.tigran@gmail.com>
diff --git a/Documentation/filesystems/btrfs.txt b/Documentation/filesystems/btrfs.txt
new file mode 100644
index 000000000..f9dad22d9
--- /dev/null
+++ b/Documentation/filesystems/btrfs.txt
@@ -0,0 +1,31 @@
+BTRFS
+=====
+
+Btrfs is a copy on write filesystem for Linux aimed at implementing advanced
+features while focusing on fault tolerance, repair and easy administration.
+Jointly developed by several companies, licensed under the GPL and open for
+contribution from anyone.
+
+The main Btrfs features include:
+
+    * Extent based file storage (2^64 max file size)
+    * Space efficient packing of small files
+    * Space efficient indexed directories
+    * Dynamic inode allocation
+    * Writable snapshots
+    * Subvolumes (separate internal filesystem roots)
+    * Object level mirroring and striping
+    * Checksums on data and metadata (multiple algorithms available)
+    * Compression
+    * Integrated multiple device support, with several raid algorithms
+    * Offline filesystem check
+    * Efficient incremental backup and FS mirroring
+    * Online filesystem defragmentation
+
+For more information please refer to the wiki
+
+  https://btrfs.wiki.kernel.org
+
+that maintains information about administration tasks, frequently asked
+questions, use cases, mount options, comprehensible changelogs, features,
+manual pages, source code repositories, contacts etc.
diff --git a/Documentation/filesystems/caching/backend-api.txt b/Documentation/filesystems/caching/backend-api.txt
new file mode 100644
index 000000000..c0bd56772
--- /dev/null
+++ b/Documentation/filesystems/caching/backend-api.txt
@@ -0,0 +1,726 @@
+			  ==========================
+			  FS-CACHE CACHE BACKEND API
+			  ==========================
+
+The FS-Cache system provides an API by which actual caches can be supplied to
+FS-Cache for it to then serve out to network filesystems and other interested
+parties.
+
+This API is declared in <linux/fscache-cache.h>.
+
+
+====================================
+INITIALISING AND REGISTERING A CACHE
+====================================
+
+To start off, a cache definition must be initialised and registered for each
+cache the backend wants to make available.  For instance, CacheFS does this in
+the fill_super() operation on mounting.
+
+The cache definition (struct fscache_cache) should be initialised by calling:
+
+	void fscache_init_cache(struct fscache_cache *cache,
+				struct fscache_cache_ops *ops,
+				const char *idfmt,
+				...);
+
+Where:
+
+ (*) "cache" is a pointer to the cache definition;
+
+ (*) "ops" is a pointer to the table of operations that the backend supports on
+     this cache; and
+
+ (*) "idfmt" is a format and printf-style arguments for constructing a label
+     for the cache.
+
+
+The cache should then be registered with FS-Cache by passing a pointer to the
+previously initialised cache definition to:
+
+	int fscache_add_cache(struct fscache_cache *cache,
+			      struct fscache_object *fsdef,
+			      const char *tagname);
+
+Two extra arguments should also be supplied:
+
+ (*) "fsdef" which should point to the object representation for the FS-Cache
+     master index in this cache.  Netfs primary index entries will be created
+     here.  FS-Cache keeps the caller's reference to the index object if
+     successful and will release it upon withdrawal of the cache.
+
+ (*) "tagname" which, if given, should be a text string naming this cache.  If
+     this is NULL, the identifier will be used instead.  For CacheFS, the
+     identifier is set to name the underlying block device and the tag can be
+     supplied by mount.
+
+This function may return -ENOMEM if it ran out of memory or -EEXIST if the tag
+is already in use.  0 will be returned on success.
+
+
+=====================
+UNREGISTERING A CACHE
+=====================
+
+A cache can be withdrawn from the system by calling this function with a
+pointer to the cache definition:
+
+	void fscache_withdraw_cache(struct fscache_cache *cache);
+
+In CacheFS's case, this is called by put_super().
+
+
+========
+SECURITY
+========
+
+The cache methods are executed one of two contexts:
+
+ (1) that of the userspace process that issued the netfs operation that caused
+     the cache method to be invoked, or
+
+ (2) that of one of the processes in the FS-Cache thread pool.
+
+In either case, this may not be an appropriate context in which to access the
+cache.
+
+The calling process's fsuid, fsgid and SELinux security identities may need to
+be masqueraded for the duration of the cache driver's access to the cache.
+This is left to the cache to handle; FS-Cache makes no effort in this regard.
+
+
+===================================
+CONTROL AND STATISTICS PRESENTATION
+===================================
+
+The cache may present data to the outside world through FS-Cache's interfaces
+in sysfs and procfs - the former for control and the latter for statistics.
+
+A sysfs directory called /sys/fs/fscache/<cachetag>/ is created if CONFIG_SYSFS
+is enabled.  This is accessible through the kobject struct fscache_cache::kobj
+and is for use by the cache as it sees fit.
+
+
+========================
+RELEVANT DATA STRUCTURES
+========================
+
+ (*) Index/Data file FS-Cache representation cookie:
+
+	struct fscache_cookie {
+		struct fscache_object_def	*def;
+		struct fscache_netfs		*netfs;
+		void				*netfs_data;
+		...
+	};
+
+     The fields that might be of use to the backend describe the object
+     definition, the netfs definition and the netfs's data for this cookie.
+     The object definition contain functions supplied by the netfs for loading
+     and matching index entries; these are required to provide some of the
+     cache operations.
+
+
+ (*) In-cache object representation:
+
+	struct fscache_object {
+		int				debug_id;
+		enum {
+			FSCACHE_OBJECT_RECYCLING,
+			...
+		}				state;
+		spinlock_t			lock
+		struct fscache_cache		*cache;
+		struct fscache_cookie		*cookie;
+		...
+	};
+
+     Structures of this type should be allocated by the cache backend and
+     passed to FS-Cache when requested by the appropriate cache operation.  In
+     the case of CacheFS, they're embedded in CacheFS's internal object
+     structures.
+
+     The debug_id is a simple integer that can be used in debugging messages
+     that refer to a particular object.  In such a case it should be printed
+     using "OBJ%x" to be consistent with FS-Cache.
+
+     Each object contains a pointer to the cookie that represents the object it
+     is backing.  An object should retired when put_object() is called if it is
+     in state FSCACHE_OBJECT_RECYCLING.  The fscache_object struct should be
+     initialised by calling fscache_object_init(object).
+
+
+ (*) FS-Cache operation record:
+
+	struct fscache_operation {
+		atomic_t		usage;
+		struct fscache_object	*object;
+		unsigned long		flags;
+	#define FSCACHE_OP_EXCLUSIVE
+		void (*processor)(struct fscache_operation *op);
+		void (*release)(struct fscache_operation *op);
+		...
+	};
+
+     FS-Cache has a pool of threads that it uses to give CPU time to the
+     various asynchronous operations that need to be done as part of driving
+     the cache.  These are represented by the above structure.  The processor
+     method is called to give the op CPU time, and the release method to get
+     rid of it when its usage count reaches 0.
+
+     An operation can be made exclusive upon an object by setting the
+     appropriate flag before enqueuing it with fscache_enqueue_operation().  If
+     an operation needs more processing time, it should be enqueued again.
+
+
+ (*) FS-Cache retrieval operation record:
+
+	struct fscache_retrieval {
+		struct fscache_operation op;
+		struct address_space	*mapping;
+		struct list_head	*to_do;
+		...
+	};
+
+     A structure of this type is allocated by FS-Cache to record retrieval and
+     allocation requests made by the netfs.  This struct is then passed to the
+     backend to do the operation.  The backend may get extra refs to it by
+     calling fscache_get_retrieval() and refs may be discarded by calling
+     fscache_put_retrieval().
+
+     A retrieval operation can be used by the backend to do retrieval work.  To
+     do this, the retrieval->op.processor method pointer should be set
+     appropriately by the backend and fscache_enqueue_retrieval() called to
+     submit it to the thread pool.  CacheFiles, for example, uses this to queue
+     page examination when it detects PG_lock being cleared.
+
+     The to_do field is an empty list available for the cache backend to use as
+     it sees fit.
+
+
+ (*) FS-Cache storage operation record:
+
+	struct fscache_storage {
+		struct fscache_operation op;
+		pgoff_t			store_limit;
+		...
+	};
+
+     A structure of this type is allocated by FS-Cache to record outstanding
+     writes to be made.  FS-Cache itself enqueues this operation and invokes
+     the write_page() method on the object at appropriate times to effect
+     storage.
+
+
+================
+CACHE OPERATIONS
+================
+
+The cache backend provides FS-Cache with a table of operations that can be
+performed on the denizens of the cache.  These are held in a structure of type:
+
+	struct fscache_cache_ops
+
+ (*) Name of cache provider [mandatory]:
+
+	const char *name
+
+     This isn't strictly an operation, but should be pointed at a string naming
+     the backend.
+
+
+ (*) Allocate a new object [mandatory]:
+
+	struct fscache_object *(*alloc_object)(struct fscache_cache *cache,
+					       struct fscache_cookie *cookie)
+
+     This method is used to allocate a cache object representation to back a
+     cookie in a particular cache.  fscache_object_init() should be called on
+     the object to initialise it prior to returning.
+
+     This function may also be used to parse the index key to be used for
+     multiple lookup calls to turn it into a more convenient form.  FS-Cache
+     will call the lookup_complete() method to allow the cache to release the
+     form once lookup is complete or aborted.
+
+
+ (*) Look up and create object [mandatory]:
+
+	void (*lookup_object)(struct fscache_object *object)
+
+     This method is used to look up an object, given that the object is already
+     allocated and attached to the cookie.  This should instantiate that object
+     in the cache if it can.
+
+     The method should call fscache_object_lookup_negative() as soon as
+     possible if it determines the object doesn't exist in the cache.  If the
+     object is found to exist and the netfs indicates that it is valid then
+     fscache_obtained_object() should be called once the object is in a
+     position to have data stored in it.  Similarly, fscache_obtained_object()
+     should also be called once a non-present object has been created.
+
+     If a lookup error occurs, fscache_object_lookup_error() should be called
+     to abort the lookup of that object.
+
+
+ (*) Release lookup data [mandatory]:
+
+	void (*lookup_complete)(struct fscache_object *object)
+
+     This method is called to ask the cache to release any resources it was
+     using to perform a lookup.
+
+
+ (*) Increment object refcount [mandatory]:
+
+	struct fscache_object *(*grab_object)(struct fscache_object *object)
+
+     This method is called to increment the reference count on an object.  It
+     may fail (for instance if the cache is being withdrawn) by returning NULL.
+     It should return the object pointer if successful.
+
+
+ (*) Lock/Unlock object [mandatory]:
+
+	void (*lock_object)(struct fscache_object *object)
+	void (*unlock_object)(struct fscache_object *object)
+
+     These methods are used to exclusively lock an object.  It must be possible
+     to schedule with the lock held, so a spinlock isn't sufficient.
+
+
+ (*) Pin/Unpin object [optional]:
+
+	int (*pin_object)(struct fscache_object *object)
+	void (*unpin_object)(struct fscache_object *object)
+
+     These methods are used to pin an object into the cache.  Once pinned an
+     object cannot be reclaimed to make space.  Return -ENOSPC if there's not
+     enough space in the cache to permit this.
+
+
+ (*) Check coherency state of an object [mandatory]:
+
+	int (*check_consistency)(struct fscache_object *object)
+
+     This method is called to have the cache check the saved auxiliary data of
+     the object against the netfs's idea of the state.  0 should be returned
+     if they're consistent and -ESTALE otherwise.  -ENOMEM and -ERESTARTSYS
+     may also be returned.
+
+ (*) Update object [mandatory]:
+
+	int (*update_object)(struct fscache_object *object)
+
+     This is called to update the index entry for the specified object.  The
+     new information should be in object->cookie->netfs_data.  This can be
+     obtained by calling object->cookie->def->get_aux()/get_attr().
+
+
+ (*) Invalidate data object [mandatory]:
+
+	int (*invalidate_object)(struct fscache_operation *op)
+
+     This is called to invalidate a data object (as pointed to by op->object).
+     All the data stored for this object should be discarded and an
+     attr_changed operation should be performed.  The caller will follow up
+     with an object update operation.
+
+     fscache_op_complete() must be called on op before returning.
+
+
+ (*) Discard object [mandatory]:
+
+	void (*drop_object)(struct fscache_object *object)
+
+     This method is called to indicate that an object has been unbound from its
+     cookie, and that the cache should release the object's resources and
+     retire it if it's in state FSCACHE_OBJECT_RECYCLING.
+
+     This method should not attempt to release any references held by the
+     caller.  The caller will invoke the put_object() method as appropriate.
+
+
+ (*) Release object reference [mandatory]:
+
+	void (*put_object)(struct fscache_object *object)
+
+     This method is used to discard a reference to an object.  The object may
+     be freed when all the references to it are released.
+
+
+ (*) Synchronise a cache [mandatory]:
+
+	void (*sync)(struct fscache_cache *cache)
+
+     This is called to ask the backend to synchronise a cache with its backing
+     device.
+
+
+ (*) Dissociate a cache [mandatory]:
+
+	void (*dissociate_pages)(struct fscache_cache *cache)
+
+     This is called to ask a cache to perform any page dissociations as part of
+     cache withdrawal.
+
+
+ (*) Notification that the attributes on a netfs file changed [mandatory]:
+
+	int (*attr_changed)(struct fscache_object *object);
+
+     This is called to indicate to the cache that certain attributes on a netfs
+     file have changed (for example the maximum size a file may reach).  The
+     cache can read these from the netfs by calling the cookie's get_attr()
+     method.
+
+     The cache may use the file size information to reserve space on the cache.
+     It should also call fscache_set_store_limit() to indicate to FS-Cache the
+     highest byte it's willing to store for an object.
+
+     This method may return -ve if an error occurred or the cache object cannot
+     be expanded.  In such a case, the object will be withdrawn from service.
+
+     This operation is run asynchronously from FS-Cache's thread pool, and
+     storage and retrieval operations from the netfs are excluded during the
+     execution of this operation.
+
+
+ (*) Reserve cache space for an object's data [optional]:
+
+	int (*reserve_space)(struct fscache_object *object, loff_t size);
+
+     This is called to request that cache space be reserved to hold the data
+     for an object and the metadata used to track it.  Zero size should be
+     taken as request to cancel a reservation.
+
+     This should return 0 if successful, -ENOSPC if there isn't enough space
+     available, or -ENOMEM or -EIO on other errors.
+
+     The reservation may exceed the current size of the object, thus permitting
+     future expansion.  If the amount of space consumed by an object would
+     exceed the reservation, it's permitted to refuse requests to allocate
+     pages, but not required.  An object may be pruned down to its reservation
+     size if larger than that already.
+
+
+ (*) Request page be read from cache [mandatory]:
+
+	int (*read_or_alloc_page)(struct fscache_retrieval *op,
+				  struct page *page,
+				  gfp_t gfp)
+
+     This is called to attempt to read a netfs page from the cache, or to
+     reserve a backing block if not.  FS-Cache will have done as much checking
+     as it can before calling, but most of the work belongs to the backend.
+
+     If there's no page in the cache, then -ENODATA should be returned if the
+     backend managed to reserve a backing block; -ENOBUFS or -ENOMEM if it
+     didn't.
+
+     If there is suitable data in the cache, then a read operation should be
+     queued and 0 returned.  When the read finishes, fscache_end_io() should be
+     called.
+
+     The fscache_mark_pages_cached() should be called for the page if any cache
+     metadata is retained.  This will indicate to the netfs that the page needs
+     explicit uncaching.  This operation takes a pagevec, thus allowing several
+     pages to be marked at once.
+
+     The retrieval record pointed to by op should be retained for each page
+     queued and released when I/O on the page has been formally ended.
+     fscache_get/put_retrieval() are available for this purpose.
+
+     The retrieval record may be used to get CPU time via the FS-Cache thread
+     pool.  If this is desired, the op->op.processor should be set to point to
+     the appropriate processing routine, and fscache_enqueue_retrieval() should
+     be called at an appropriate point to request CPU time.  For instance, the
+     retrieval routine could be enqueued upon the completion of a disk read.
+     The to_do field in the retrieval record is provided to aid in this.
+
+     If an I/O error occurs, fscache_io_error() should be called and -ENOBUFS
+     returned if possible or fscache_end_io() called with a suitable error
+     code.
+
+     fscache_put_retrieval() should be called after a page or pages are dealt
+     with.  This will complete the operation when all pages are dealt with.
+
+
+ (*) Request pages be read from cache [mandatory]:
+
+	int (*read_or_alloc_pages)(struct fscache_retrieval *op,
+				   struct list_head *pages,
+				   unsigned *nr_pages,
+				   gfp_t gfp)
+
+     This is like the read_or_alloc_page() method, except it is handed a list
+     of pages instead of one page.  Any pages on which a read operation is
+     started must be added to the page cache for the specified mapping and also
+     to the LRU.  Such pages must also be removed from the pages list and
+     *nr_pages decremented per page.
+
+     If there was an error such as -ENOMEM, then that should be returned; else
+     if one or more pages couldn't be read or allocated, then -ENOBUFS should
+     be returned; else if one or more pages couldn't be read, then -ENODATA
+     should be returned.  If all the pages are dispatched then 0 should be
+     returned.
+
+
+ (*) Request page be allocated in the cache [mandatory]:
+
+	int (*allocate_page)(struct fscache_retrieval *op,
+			     struct page *page,
+			     gfp_t gfp)
+
+     This is like the read_or_alloc_page() method, except that it shouldn't
+     read from the cache, even if there's data there that could be retrieved.
+     It should, however, set up any internal metadata required such that
+     the write_page() method can write to the cache.
+
+     If there's no backing block available, then -ENOBUFS should be returned
+     (or -ENOMEM if there were other problems).  If a block is successfully
+     allocated, then the netfs page should be marked and 0 returned.
+
+
+ (*) Request pages be allocated in the cache [mandatory]:
+
+	int (*allocate_pages)(struct fscache_retrieval *op,
+			      struct list_head *pages,
+			      unsigned *nr_pages,
+			      gfp_t gfp)
+
+     This is an multiple page version of the allocate_page() method.  pages and
+     nr_pages should be treated as for the read_or_alloc_pages() method.
+
+
+ (*) Request page be written to cache [mandatory]:
+
+	int (*write_page)(struct fscache_storage *op,
+			  struct page *page);
+
+     This is called to write from a page on which there was a previously
+     successful read_or_alloc_page() call or similar.  FS-Cache filters out
+     pages that don't have mappings.
+
+     This method is called asynchronously from the FS-Cache thread pool.  It is
+     not required to actually store anything, provided -ENODATA is then
+     returned to the next read of this page.
+
+     If an error occurred, then a negative error code should be returned,
+     otherwise zero should be returned.  FS-Cache will take appropriate action
+     in response to an error, such as withdrawing this object.
+
+     If this method returns success then FS-Cache will inform the netfs
+     appropriately.
+
+
+ (*) Discard retained per-page metadata [mandatory]:
+
+	void (*uncache_page)(struct fscache_object *object, struct page *page)
+
+     This is called when a netfs page is being evicted from the pagecache.  The
+     cache backend should tear down any internal representation or tracking it
+     maintains for this page.
+
+
+==================
+FS-CACHE UTILITIES
+==================
+
+FS-Cache provides some utilities that a cache backend may make use of:
+
+ (*) Note occurrence of an I/O error in a cache:
+
+	void fscache_io_error(struct fscache_cache *cache)
+
+     This tells FS-Cache that an I/O error occurred in the cache.  After this
+     has been called, only resource dissociation operations (object and page
+     release) will be passed from the netfs to the cache backend for the
+     specified cache.
+
+     This does not actually withdraw the cache.  That must be done separately.
+
+
+ (*) Invoke the retrieval I/O completion function:
+
+	void fscache_end_io(struct fscache_retrieval *op, struct page *page,
+			    int error);
+
+     This is called to note the end of an attempt to retrieve a page.  The
+     error value should be 0 if successful and an error otherwise.
+
+
+ (*) Record that one or more pages being retrieved or allocated have been dealt
+     with:
+
+	void fscache_retrieval_complete(struct fscache_retrieval *op,
+					int n_pages);
+
+     This is called to record the fact that one or more pages have been dealt
+     with and are no longer the concern of this operation.  When the number of
+     pages remaining in the operation reaches 0, the operation will be
+     completed.
+
+
+ (*) Record operation completion:
+
+	void fscache_op_complete(struct fscache_operation *op);
+
+     This is called to record the completion of an operation.  This deducts
+     this operation from the parent object's run state, potentially permitting
+     one or more pending operations to start running.
+
+
+ (*) Set highest store limit:
+
+	void fscache_set_store_limit(struct fscache_object *object,
+				     loff_t i_size);
+
+     This sets the limit FS-Cache imposes on the highest byte it's willing to
+     try and store for a netfs.  Any page over this limit is automatically
+     rejected by fscache_read_alloc_page() and co with -ENOBUFS.
+
+
+ (*) Mark pages as being cached:
+
+	void fscache_mark_pages_cached(struct fscache_retrieval *op,
+				       struct pagevec *pagevec);
+
+     This marks a set of pages as being cached.  After this has been called,
+     the netfs must call fscache_uncache_page() to unmark the pages.
+
+
+ (*) Perform coherency check on an object:
+
+	enum fscache_checkaux fscache_check_aux(struct fscache_object *object,
+						const void *data,
+						uint16_t datalen);
+
+     This asks the netfs to perform a coherency check on an object that has
+     just been looked up.  The cookie attached to the object will determine the
+     netfs to use.  data and datalen should specify where the auxiliary data
+     retrieved from the cache can be found.
+
+     One of three values will be returned:
+
+	(*) FSCACHE_CHECKAUX_OKAY
+
+	    The coherency data indicates the object is valid as is.
+
+	(*) FSCACHE_CHECKAUX_NEEDS_UPDATE
+
+	    The coherency data needs updating, but otherwise the object is
+	    valid.
+
+	(*) FSCACHE_CHECKAUX_OBSOLETE
+
+	    The coherency data indicates that the object is obsolete and should
+	    be discarded.
+
+
+ (*) Initialise a freshly allocated object:
+
+	void fscache_object_init(struct fscache_object *object);
+
+     This initialises all the fields in an object representation.
+
+
+ (*) Indicate the destruction of an object:
+
+	void fscache_object_destroyed(struct fscache_cache *cache);
+
+     This must be called to inform FS-Cache that an object that belonged to a
+     cache has been destroyed and deallocated.  This will allow continuation
+     of the cache withdrawal process when it is stopped pending destruction of
+     all the objects.
+
+
+ (*) Indicate negative lookup on an object:
+
+	void fscache_object_lookup_negative(struct fscache_object *object);
+
+     This is called to indicate to FS-Cache that a lookup process for an object
+     found a negative result.
+
+     This changes the state of an object to permit reads pending on lookup
+     completion to go off and start fetching data from the netfs server as it's
+     known at this point that there can't be any data in the cache.
+
+     This may be called multiple times on an object.  Only the first call is
+     significant - all subsequent calls are ignored.
+
+
+ (*) Indicate an object has been obtained:
+
+	void fscache_obtained_object(struct fscache_object *object);
+
+     This is called to indicate to FS-Cache that a lookup process for an object
+     produced a positive result, or that an object was created.  This should
+     only be called once for any particular object.
+
+     This changes the state of an object to indicate:
+
+	(1) if no call to fscache_object_lookup_negative() has been made on
+	    this object, that there may be data available, and that reads can
+	    now go and look for it; and
+
+        (2) that writes may now proceed against this object.
+
+
+ (*) Indicate that object lookup failed:
+
+	void fscache_object_lookup_error(struct fscache_object *object);
+
+     This marks an object as having encountered a fatal error (usually EIO)
+     and causes it to move into a state whereby it will be withdrawn as soon
+     as possible.
+
+
+ (*) Indicate that a stale object was found and discarded:
+
+	void fscache_object_retrying_stale(struct fscache_object *object);
+
+     This is called to indicate that the lookup procedure found an object in
+     the cache that the netfs decided was stale.  The object has been
+     discarded from the cache and the lookup will be performed again.
+
+
+ (*) Indicate that the caching backend killed an object:
+
+	void fscache_object_mark_killed(struct fscache_object *object,
+					enum fscache_why_object_killed why);
+
+     This is called to indicate that the cache backend preemptively killed an
+     object.  The why parameter should be set to indicate the reason:
+
+	FSCACHE_OBJECT_IS_STALE - the object was stale and needs discarding.
+	FSCACHE_OBJECT_NO_SPACE - there was insufficient cache space
+	FSCACHE_OBJECT_WAS_RETIRED - the object was retired when relinquished.
+	FSCACHE_OBJECT_WAS_CULLED - the object was culled to make space.
+
+
+ (*) Get and release references on a retrieval record:
+
+	void fscache_get_retrieval(struct fscache_retrieval *op);
+	void fscache_put_retrieval(struct fscache_retrieval *op);
+
+     These two functions are used to retain a retrieval record whilst doing
+     asynchronous data retrieval and block allocation.
+
+
+ (*) Enqueue a retrieval record for processing.
+
+	void fscache_enqueue_retrieval(struct fscache_retrieval *op);
+
+     This enqueues a retrieval record for processing by the FS-Cache thread
+     pool.  One of the threads in the pool will invoke the retrieval record's
+     op->op.processor callback function.  This function may be called from
+     within the callback function.
+
+
+ (*) List of object state names:
+
+	const char *fscache_object_states[];
+
+     For debugging purposes, this may be used to turn the state that an object
+     is in into a text string for display purposes.
diff --git a/Documentation/filesystems/caching/cachefiles.txt b/Documentation/filesystems/caching/cachefiles.txt
new file mode 100644
index 000000000..748a1ae49
--- /dev/null
+++ b/Documentation/filesystems/caching/cachefiles.txt
@@ -0,0 +1,501 @@
+	       ===============================================
+	       CacheFiles: CACHE ON ALREADY MOUNTED FILESYSTEM
+	       ===============================================
+
+Contents:
+
+ (*) Overview.
+
+ (*) Requirements.
+
+ (*) Configuration.
+
+ (*) Starting the cache.
+
+ (*) Things to avoid.
+
+ (*) Cache culling.
+
+ (*) Cache structure.
+
+ (*) Security model and SELinux.
+
+ (*) A note on security.
+
+ (*) Statistical information.
+
+ (*) Debugging.
+
+
+========
+OVERVIEW
+========
+
+CacheFiles is a caching backend that's meant to use as a cache a directory on
+an already mounted filesystem of a local type (such as Ext3).
+
+CacheFiles uses a userspace daemon to do some of the cache management - such as
+reaping stale nodes and culling.  This is called cachefilesd and lives in
+/sbin.
+
+The filesystem and data integrity of the cache are only as good as those of the
+filesystem providing the backing services.  Note that CacheFiles does not
+attempt to journal anything since the journalling interfaces of the various
+filesystems are very specific in nature.
+
+CacheFiles creates a misc character device - "/dev/cachefiles" - that is used
+to communication with the daemon.  Only one thing may have this open at once,
+and whilst it is open, a cache is at least partially in existence.  The daemon
+opens this and sends commands down it to control the cache.
+
+CacheFiles is currently limited to a single cache.
+
+CacheFiles attempts to maintain at least a certain percentage of free space on
+the filesystem, shrinking the cache by culling the objects it contains to make
+space if necessary - see the "Cache Culling" section.  This means it can be
+placed on the same medium as a live set of data, and will expand to make use of
+spare space and automatically contract when the set of data requires more
+space.
+
+
+============
+REQUIREMENTS
+============
+
+The use of CacheFiles and its daemon requires the following features to be
+available in the system and in the cache filesystem:
+
+	- dnotify.
+
+	- extended attributes (xattrs).
+
+	- openat() and friends.
+
+	- bmap() support on files in the filesystem (FIBMAP ioctl).
+
+	- The use of bmap() to detect a partial page at the end of the file.
+
+It is strongly recommended that the "dir_index" option is enabled on Ext3
+filesystems being used as a cache.
+
+
+=============
+CONFIGURATION
+=============
+
+The cache is configured by a script in /etc/cachefilesd.conf.  These commands
+set up cache ready for use.  The following script commands are available:
+
+ (*) brun <N>%
+ (*) bcull <N>%
+ (*) bstop <N>%
+ (*) frun <N>%
+ (*) fcull <N>%
+ (*) fstop <N>%
+
+	Configure the culling limits.  Optional.  See the section on culling
+	The defaults are 7% (run), 5% (cull) and 1% (stop) respectively.
+
+	The commands beginning with a 'b' are file space (block) limits, those
+	beginning with an 'f' are file count limits.
+
+ (*) dir <path>
+
+	Specify the directory containing the root of the cache.  Mandatory.
+
+ (*) tag <name>
+
+	Specify a tag to FS-Cache to use in distinguishing multiple caches.
+	Optional.  The default is "CacheFiles".
+
+ (*) debug <mask>
+
+	Specify a numeric bitmask to control debugging in the kernel module.
+	Optional.  The default is zero (all off).  The following values can be
+	OR'd into the mask to collect various information:
+
+		1	Turn on trace of function entry (_enter() macros)
+		2	Turn on trace of function exit (_leave() macros)
+		4	Turn on trace of internal debug points (_debug())
+
+	This mask can also be set through sysfs, eg:
+
+		echo 5 >/sys/modules/cachefiles/parameters/debug
+
+
+==================
+STARTING THE CACHE
+==================
+
+The cache is started by running the daemon.  The daemon opens the cache device,
+configures the cache and tells it to begin caching.  At that point the cache
+binds to fscache and the cache becomes live.
+
+The daemon is run as follows:
+
+	/sbin/cachefilesd [-d]* [-s] [-n] [-f <configfile>]
+
+The flags are:
+
+ (*) -d
+
+	Increase the debugging level.  This can be specified multiple times and
+	is cumulative with itself.
+
+ (*) -s
+
+	Send messages to stderr instead of syslog.
+
+ (*) -n
+
+	Don't daemonise and go into background.
+
+ (*) -f <configfile>
+
+	Use an alternative configuration file rather than the default one.
+
+
+===============
+THINGS TO AVOID
+===============
+
+Do not mount other things within the cache as this will cause problems.  The
+kernel module contains its own very cut-down path walking facility that ignores
+mountpoints, but the daemon can't avoid them.
+
+Do not create, rename or unlink files and directories in the cache whilst the
+cache is active, as this may cause the state to become uncertain.
+
+Renaming files in the cache might make objects appear to be other objects (the
+filename is part of the lookup key).
+
+Do not change or remove the extended attributes attached to cache files by the
+cache as this will cause the cache state management to get confused.
+
+Do not create files or directories in the cache, lest the cache get confused or
+serve incorrect data.
+
+Do not chmod files in the cache.  The module creates things with minimal
+permissions to prevent random users being able to access them directly.
+
+
+=============
+CACHE CULLING
+=============
+
+The cache may need culling occasionally to make space.  This involves
+discarding objects from the cache that have been used less recently than
+anything else.  Culling is based on the access time of data objects.  Empty
+directories are culled if not in use.
+
+Cache culling is done on the basis of the percentage of blocks and the
+percentage of files available in the underlying filesystem.  There are six
+"limits":
+
+ (*) brun
+ (*) frun
+
+     If the amount of free space and the number of available files in the cache
+     rises above both these limits, then culling is turned off.
+
+ (*) bcull
+ (*) fcull
+
+     If the amount of available space or the number of available files in the
+     cache falls below either of these limits, then culling is started.
+
+ (*) bstop
+ (*) fstop
+
+     If the amount of available space or the number of available files in the
+     cache falls below either of these limits, then no further allocation of
+     disk space or files is permitted until culling has raised things above
+     these limits again.
+
+These must be configured thusly:
+
+	0 <= bstop < bcull < brun < 100
+	0 <= fstop < fcull < frun < 100
+
+Note that these are percentages of available space and available files, and do
+_not_ appear as 100 minus the percentage displayed by the "df" program.
+
+The userspace daemon scans the cache to build up a table of cullable objects.
+These are then culled in least recently used order.  A new scan of the cache is
+started as soon as space is made in the table.  Objects will be skipped if
+their atimes have changed or if the kernel module says it is still using them.
+
+
+===============
+CACHE STRUCTURE
+===============
+
+The CacheFiles module will create two directories in the directory it was
+given:
+
+ (*) cache/
+
+ (*) graveyard/
+
+The active cache objects all reside in the first directory.  The CacheFiles
+kernel module moves any retired or culled objects that it can't simply unlink
+to the graveyard from which the daemon will actually delete them.
+
+The daemon uses dnotify to monitor the graveyard directory, and will delete
+anything that appears therein.
+
+
+The module represents index objects as directories with the filename "I..." or
+"J...".  Note that the "cache/" directory is itself a special index.
+
+Data objects are represented as files if they have no children, or directories
+if they do.  Their filenames all begin "D..." or "E...".  If represented as a
+directory, data objects will have a file in the directory called "data" that
+actually holds the data.
+
+Special objects are similar to data objects, except their filenames begin
+"S..." or "T...".
+
+
+If an object has children, then it will be represented as a directory.
+Immediately in the representative directory are a collection of directories
+named for hash values of the child object keys with an '@' prepended.  Into
+this directory, if possible, will be placed the representations of the child
+objects:
+
+	INDEX     INDEX      INDEX                             DATA FILES
+	========= ========== ================================= ================
+	cache/@4a/I03nfs/@30/Ji000000000000000--fHg8hi8400
+	cache/@4a/I03nfs/@30/Ji000000000000000--fHg8hi8400/@75/Es0g000w...DB1ry
+	cache/@4a/I03nfs/@30/Ji000000000000000--fHg8hi8400/@75/Es0g000w...N22ry
+	cache/@4a/I03nfs/@30/Ji000000000000000--fHg8hi8400/@75/Es0g000w...FP1ry
+
+
+If the key is so long that it exceeds NAME_MAX with the decorations added on to
+it, then it will be cut into pieces, the first few of which will be used to
+make a nest of directories, and the last one of which will be the objects
+inside the last directory.  The names of the intermediate directories will have
+'+' prepended:
+
+	J1223/@23/+xy...z/+kl...m/Epqr
+
+
+Note that keys are raw data, and not only may they exceed NAME_MAX in size,
+they may also contain things like '/' and NUL characters, and so they may not
+be suitable for turning directly into a filename.
+
+To handle this, CacheFiles will use a suitably printable filename directly and
+"base-64" encode ones that aren't directly suitable.  The two versions of
+object filenames indicate the encoding:
+
+	OBJECT TYPE	PRINTABLE	ENCODED
+	===============	===============	===============
+	Index		"I..."		"J..."
+	Data		"D..."		"E..."
+	Special		"S..."		"T..."
+
+Intermediate directories are always "@" or "+" as appropriate.
+
+
+Each object in the cache has an extended attribute label that holds the object
+type ID (required to distinguish special objects) and the auxiliary data from
+the netfs.  The latter is used to detect stale objects in the cache and update
+or retire them.
+
+
+Note that CacheFiles will erase from the cache any file it doesn't recognise or
+any file of an incorrect type (such as a FIFO file or a device file).
+
+
+==========================
+SECURITY MODEL AND SELINUX
+==========================
+
+CacheFiles is implemented to deal properly with the LSM security features of
+the Linux kernel and the SELinux facility.
+
+One of the problems that CacheFiles faces is that it is generally acting on
+behalf of a process, and running in that process's context, and that includes a
+security context that is not appropriate for accessing the cache - either
+because the files in the cache are inaccessible to that process, or because if
+the process creates a file in the cache, that file may be inaccessible to other
+processes.
+
+The way CacheFiles works is to temporarily change the security context (fsuid,
+fsgid and actor security label) that the process acts as - without changing the
+security context of the process when it the target of an operation performed by
+some other process (so signalling and suchlike still work correctly).
+
+
+When the CacheFiles module is asked to bind to its cache, it:
+
+ (1) Finds the security label attached to the root cache directory and uses
+     that as the security label with which it will create files.  By default,
+     this is:
+
+	cachefiles_var_t
+
+ (2) Finds the security label of the process which issued the bind request
+     (presumed to be the cachefilesd daemon), which by default will be:
+
+	cachefilesd_t
+
+     and asks LSM to supply a security ID as which it should act given the
+     daemon's label.  By default, this will be:
+
+	cachefiles_kernel_t
+
+     SELinux transitions the daemon's security ID to the module's security ID
+     based on a rule of this form in the policy.
+
+	type_transition <daemon's-ID> kernel_t : process <module's-ID>;
+
+     For instance:
+
+	type_transition cachefilesd_t kernel_t : process cachefiles_kernel_t;
+
+
+The module's security ID gives it permission to create, move and remove files
+and directories in the cache, to find and access directories and files in the
+cache, to set and access extended attributes on cache objects, and to read and
+write files in the cache.
+
+The daemon's security ID gives it only a very restricted set of permissions: it
+may scan directories, stat files and erase files and directories.  It may
+not read or write files in the cache, and so it is precluded from accessing the
+data cached therein; nor is it permitted to create new files in the cache.
+
+
+There are policy source files available in:
+
+	http://people.redhat.com/~dhowells/fscache/cachefilesd-0.8.tar.bz2
+
+and later versions.  In that tarball, see the files:
+
+	cachefilesd.te
+	cachefilesd.fc
+	cachefilesd.if
+
+They are built and installed directly by the RPM.
+
+If a non-RPM based system is being used, then copy the above files to their own
+directory and run:
+
+	make -f /usr/share/selinux/devel/Makefile
+	semodule -i cachefilesd.pp
+
+You will need checkpolicy and selinux-policy-devel installed prior to the
+build.
+
+
+By default, the cache is located in /var/fscache, but if it is desirable that
+it should be elsewhere, than either the above policy files must be altered, or
+an auxiliary policy must be installed to label the alternate location of the
+cache.
+
+For instructions on how to add an auxiliary policy to enable the cache to be
+located elsewhere when SELinux is in enforcing mode, please see:
+
+	/usr/share/doc/cachefilesd-*/move-cache.txt
+
+When the cachefilesd rpm is installed; alternatively, the document can be found
+in the sources.
+
+
+==================
+A NOTE ON SECURITY
+==================
+
+CacheFiles makes use of the split security in the task_struct.  It allocates
+its own task_security structure, and redirects current->cred to point to it
+when it acts on behalf of another process, in that process's context.
+
+The reason it does this is that it calls vfs_mkdir() and suchlike rather than
+bypassing security and calling inode ops directly.  Therefore the VFS and LSM
+may deny the CacheFiles access to the cache data because under some
+circumstances the caching code is running in the security context of whatever
+process issued the original syscall on the netfs.
+
+Furthermore, should CacheFiles create a file or directory, the security
+parameters with that object is created (UID, GID, security label) would be
+derived from that process that issued the system call, thus potentially
+preventing other processes from accessing the cache - including CacheFiles's
+cache management daemon (cachefilesd).
+
+What is required is to temporarily override the security of the process that
+issued the system call.  We can't, however, just do an in-place change of the
+security data as that affects the process as an object, not just as a subject.
+This means it may lose signals or ptrace events for example, and affects what
+the process looks like in /proc.
+
+So CacheFiles makes use of a logical split in the security between the
+objective security (task->real_cred) and the subjective security (task->cred).
+The objective security holds the intrinsic security properties of a process and
+is never overridden.  This is what appears in /proc, and is what is used when a
+process is the target of an operation by some other process (SIGKILL for
+example).
+
+The subjective security holds the active security properties of a process, and
+may be overridden.  This is not seen externally, and is used whan a process
+acts upon another object, for example SIGKILLing another process or opening a
+file.
+
+LSM hooks exist that allow SELinux (or Smack or whatever) to reject a request
+for CacheFiles to run in a context of a specific security label, or to create
+files and directories with another security label.
+
+
+=======================
+STATISTICAL INFORMATION
+=======================
+
+If FS-Cache is compiled with the following option enabled:
+
+	CONFIG_CACHEFILES_HISTOGRAM=y
+
+then it will gather certain statistics and display them through a proc file.
+
+ (*) /proc/fs/cachefiles/histogram
+
+	cat /proc/fs/cachefiles/histogram
+	JIFS  SECS  LOOKUPS   MKDIRS    CREATES
+	===== ===== ========= ========= =========
+
+     This shows the breakdown of the number of times each amount of time
+     between 0 jiffies and HZ-1 jiffies a variety of tasks took to run.  The
+     columns are as follows:
+
+	COLUMN		TIME MEASUREMENT
+	=======		=======================================================
+	LOOKUPS		Length of time to perform a lookup on the backing fs
+	MKDIRS		Length of time to perform a mkdir on the backing fs
+	CREATES		Length of time to perform a create on the backing fs
+
+     Each row shows the number of events that took a particular range of times.
+     Each step is 1 jiffy in size.  The JIFS column indicates the particular
+     jiffy range covered, and the SECS field the equivalent number of seconds.
+
+
+=========
+DEBUGGING
+=========
+
+If CONFIG_CACHEFILES_DEBUG is enabled, the CacheFiles facility can have runtime
+debugging enabled by adjusting the value in:
+
+	/sys/module/cachefiles/parameters/debug
+
+This is a bitmask of debugging streams to enable:
+
+	BIT	VALUE	STREAM				POINT
+	=======	=======	===============================	=======================
+	0	1	General				Function entry trace
+	1	2					Function exit trace
+	2	4					General
+
+The appropriate set of values should be OR'd together and the result written to
+the control file.  For example:
+
+	echo $((1|4|8)) >/sys/module/cachefiles/parameters/debug
+
+will turn on all function entry debugging.
diff --git a/Documentation/filesystems/caching/fscache.txt b/Documentation/filesystems/caching/fscache.txt
new file mode 100644
index 000000000..50f0a5757
--- /dev/null
+++ b/Documentation/filesystems/caching/fscache.txt
@@ -0,0 +1,448 @@
+			  ==========================
+			  General Filesystem Caching
+			  ==========================
+
+========
+OVERVIEW
+========
+
+This facility is a general purpose cache for network filesystems, though it
+could be used for caching other things such as ISO9660 filesystems too.
+
+FS-Cache mediates between cache backends (such as CacheFS) and network
+filesystems:
+
+	+---------+
+	|         |                        +--------------+
+	|   NFS   |--+                     |              |
+	|         |  |                 +-->|   CacheFS    |
+	+---------+  |   +----------+  |   |  /dev/hda5   |
+	             |   |          |  |   +--------------+
+	+---------+  +-->|          |  |
+	|         |      |          |--+
+	|   AFS   |----->| FS-Cache |
+	|         |      |          |--+
+	+---------+  +-->|          |  |
+	             |   |          |  |   +--------------+
+	+---------+  |   +----------+  |   |              |
+	|         |  |                 +-->|  CacheFiles  |
+	|  ISOFS  |--+                     |  /var/cache  |
+	|         |                        +--------------+
+	+---------+
+
+Or to look at it another way, FS-Cache is a module that provides a caching
+facility to a network filesystem such that the cache is transparent to the
+user:
+
+	+---------+
+	|         |
+	| Server  |
+	|         |
+	+---------+
+	     |                  NETWORK
+	~~~~~|~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+	     |
+	     |           +----------+
+	     V           |          |
+	+---------+      |          |
+	|         |      |          |
+	|   NFS   |----->| FS-Cache |
+	|         |      |          |--+
+	+---------+      |          |  |   +--------------+   +--------------+
+	     |           |          |  |   |              |   |              |
+	     V           +----------+  +-->|  CacheFiles  |-->|  Ext3        |
+	+---------+                        |  /var/cache  |   |  /dev/sda6   |
+	|         |                        +--------------+   +--------------+
+	|   VFS   |                                ^                     ^
+	|         |                                |                     |
+	+---------+                                +--------------+      |
+	     |                  KERNEL SPACE                      |      |
+	~~~~~|~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~|~~~~~~|~~~~
+	     |                  USER SPACE                        |      |
+	     V                                                    |      |
+	+---------+                                           +--------------+
+	|         |                                           |              |
+	| Process |                                           | cachefilesd  |
+	|         |                                           |              |
+	+---------+                                           +--------------+
+
+
+FS-Cache does not follow the idea of completely loading every netfs file
+opened in its entirety into a cache before permitting it to be accessed and
+then serving the pages out of that cache rather than the netfs inode because:
+
+ (1) It must be practical to operate without a cache.
+
+ (2) The size of any accessible file must not be limited to the size of the
+     cache.
+
+ (3) The combined size of all opened files (this includes mapped libraries)
+     must not be limited to the size of the cache.
+
+ (4) The user should not be forced to download an entire file just to do a
+     one-off access of a small portion of it (such as might be done with the
+     "file" program).
+
+It instead serves the cache out in PAGE_SIZE chunks as and when requested by
+the netfs('s) using it.
+
+
+FS-Cache provides the following facilities:
+
+ (1) More than one cache can be used at once.  Caches can be selected
+     explicitly by use of tags.
+
+ (2) Caches can be added / removed at any time.
+
+ (3) The netfs is provided with an interface that allows either party to
+     withdraw caching facilities from a file (required for (2)).
+
+ (4) The interface to the netfs returns as few errors as possible, preferring
+     rather to let the netfs remain oblivious.
+
+ (5) Cookies are used to represent indices, files and other objects to the
+     netfs.  The simplest cookie is just a NULL pointer - indicating nothing
+     cached there.
+
+ (6) The netfs is allowed to propose - dynamically - any index hierarchy it
+     desires, though it must be aware that the index search function is
+     recursive, stack space is limited, and indices can only be children of
+     indices.
+
+ (7) Data I/O is done direct to and from the netfs's pages.  The netfs
+     indicates that page A is at index B of the data-file represented by cookie
+     C, and that it should be read or written.  The cache backend may or may
+     not start I/O on that page, but if it does, a netfs callback will be
+     invoked to indicate completion.  The I/O may be either synchronous or
+     asynchronous.
+
+ (8) Cookies can be "retired" upon release.  At this point FS-Cache will mark
+     them as obsolete and the index hierarchy rooted at that point will get
+     recycled.
+
+ (9) The netfs provides a "match" function for index searches.  In addition to
+     saying whether a match was made or not, this can also specify that an
+     entry should be updated or deleted.
+
+(10) As much as possible is done asynchronously.
+
+
+FS-Cache maintains a virtual indexing tree in which all indices, files, objects
+and pages are kept.  Bits of this tree may actually reside in one or more
+caches.
+
+                                           FSDEF
+                                             |
+                        +------------------------------------+
+                        |                                    |
+                       NFS                                  AFS
+                        |                                    |
+           +--------------------------+                +-----------+
+           |                          |                |           |
+        homedir                     mirror          afs.org   redhat.com
+           |                          |                            |
+     +------------+           +---------------+              +----------+
+     |            |           |               |              |          |
+   00001        00002       00007           00125        vol00001   vol00002
+     |            |           |               |                         |
+ +---+---+     +-----+      +---+      +------+------+            +-----+----+
+ |   |   |     |     |      |   |      |      |      |            |     |    |
+PG0 PG1 PG2   PG0  XATTR   PG0 PG1   DIRENT DIRENT DIRENT        R/W   R/O  Bak
+                     |                                            |
+                    PG0                                       +-------+
+                                                              |       |
+                                                            00001   00003
+                                                              |
+                                                          +---+---+
+                                                          |   |   |
+                                                         PG0 PG1 PG2
+
+In the example above, you can see two netfs's being backed: NFS and AFS.  These
+have different index hierarchies:
+
+ (*) The NFS primary index contains per-server indices.  Each server index is
+     indexed by NFS file handles to get data file objects.  Each data file
+     objects can have an array of pages, but may also have further child
+     objects, such as extended attributes and directory entries.  Extended
+     attribute objects themselves have page-array contents.
+
+ (*) The AFS primary index contains per-cell indices.  Each cell index contains
+     per-logical-volume indices.  Each of volume index contains up to three
+     indices for the read-write, read-only and backup mirrors of those volumes.
+     Each of these contains vnode data file objects, each of which contains an
+     array of pages.
+
+The very top index is the FS-Cache master index in which individual netfs's
+have entries.
+
+Any index object may reside in more than one cache, provided it only has index
+children.  Any index with non-index object children will be assumed to only
+reside in one cache.
+
+
+The netfs API to FS-Cache can be found in:
+
+	Documentation/filesystems/caching/netfs-api.txt
+
+The cache backend API to FS-Cache can be found in:
+
+	Documentation/filesystems/caching/backend-api.txt
+
+A description of the internal representations and object state machine can be
+found in:
+
+	Documentation/filesystems/caching/object.txt
+
+
+=======================
+STATISTICAL INFORMATION
+=======================
+
+If FS-Cache is compiled with the following options enabled:
+
+	CONFIG_FSCACHE_STATS=y
+	CONFIG_FSCACHE_HISTOGRAM=y
+
+then it will gather certain statistics and display them through a number of
+proc files.
+
+ (*) /proc/fs/fscache/stats
+
+     This shows counts of a number of events that can happen in FS-Cache:
+
+	CLASS	EVENT	MEANING
+	=======	=======	=======================================================
+	Cookies	idx=N	Number of index cookies allocated
+		dat=N	Number of data storage cookies allocated
+		spc=N	Number of special cookies allocated
+	Objects	alc=N	Number of objects allocated
+		nal=N	Number of object allocation failures
+		avl=N	Number of objects that reached the available state
+		ded=N	Number of objects that reached the dead state
+	ChkAux	non=N	Number of objects that didn't have a coherency check
+		ok=N	Number of objects that passed a coherency check
+		upd=N	Number of objects that needed a coherency data update
+		obs=N	Number of objects that were declared obsolete
+	Pages	mrk=N	Number of pages marked as being cached
+		unc=N	Number of uncache page requests seen
+	Acquire	n=N	Number of acquire cookie requests seen
+		nul=N	Number of acq reqs given a NULL parent
+		noc=N	Number of acq reqs rejected due to no cache available
+		ok=N	Number of acq reqs succeeded
+		nbf=N	Number of acq reqs rejected due to error
+		oom=N	Number of acq reqs failed on ENOMEM
+	Lookups	n=N	Number of lookup calls made on cache backends
+		neg=N	Number of negative lookups made
+		pos=N	Number of positive lookups made
+		crt=N	Number of objects created by lookup
+		tmo=N	Number of lookups timed out and requeued
+	Updates	n=N	Number of update cookie requests seen
+		nul=N	Number of upd reqs given a NULL parent
+		run=N	Number of upd reqs granted CPU time
+	Relinqs	n=N	Number of relinquish cookie requests seen
+		nul=N	Number of rlq reqs given a NULL parent
+		wcr=N	Number of rlq reqs waited on completion of creation
+	AttrChg	n=N	Number of attribute changed requests seen
+		ok=N	Number of attr changed requests queued
+		nbf=N	Number of attr changed rejected -ENOBUFS
+		oom=N	Number of attr changed failed -ENOMEM
+		run=N	Number of attr changed ops given CPU time
+	Allocs	n=N	Number of allocation requests seen
+		ok=N	Number of successful alloc reqs
+		wt=N	Number of alloc reqs that waited on lookup completion
+		nbf=N	Number of alloc reqs rejected -ENOBUFS
+		int=N	Number of alloc reqs aborted -ERESTARTSYS
+		ops=N	Number of alloc reqs submitted
+		owt=N	Number of alloc reqs waited for CPU time
+		abt=N	Number of alloc reqs aborted due to object death
+	Retrvls	n=N	Number of retrieval (read) requests seen
+		ok=N	Number of successful retr reqs
+		wt=N	Number of retr reqs that waited on lookup completion
+		nod=N	Number of retr reqs returned -ENODATA
+		nbf=N	Number of retr reqs rejected -ENOBUFS
+		int=N	Number of retr reqs aborted -ERESTARTSYS
+		oom=N	Number of retr reqs failed -ENOMEM
+		ops=N	Number of retr reqs submitted
+		owt=N	Number of retr reqs waited for CPU time
+		abt=N	Number of retr reqs aborted due to object death
+	Stores	n=N	Number of storage (write) requests seen
+		ok=N	Number of successful store reqs
+		agn=N	Number of store reqs on a page already pending storage
+		nbf=N	Number of store reqs rejected -ENOBUFS
+		oom=N	Number of store reqs failed -ENOMEM
+		ops=N	Number of store reqs submitted
+		run=N	Number of store reqs granted CPU time
+		pgs=N	Number of pages given store req processing time
+		rxd=N	Number of store reqs deleted from tracking tree
+		olm=N	Number of store reqs over store limit
+	VmScan	nos=N	Number of release reqs against pages with no pending store
+		gon=N	Number of release reqs against pages stored by time lock granted
+		bsy=N	Number of release reqs ignored due to in-progress store
+		can=N	Number of page stores cancelled due to release req
+	Ops	pend=N	Number of times async ops added to pending queues
+		run=N	Number of times async ops given CPU time
+		enq=N	Number of times async ops queued for processing
+		can=N	Number of async ops cancelled
+		rej=N	Number of async ops rejected due to object lookup/create failure
+		ini=N	Number of async ops initialised
+		dfr=N	Number of async ops queued for deferred release
+		rel=N	Number of async ops released (should equal ini=N when idle)
+		gc=N	Number of deferred-release async ops garbage collected
+	CacheOp	alo=N	Number of in-progress alloc_object() cache ops
+		luo=N	Number of in-progress lookup_object() cache ops
+		luc=N	Number of in-progress lookup_complete() cache ops
+		gro=N	Number of in-progress grab_object() cache ops
+		upo=N	Number of in-progress update_object() cache ops
+		dro=N	Number of in-progress drop_object() cache ops
+		pto=N	Number of in-progress put_object() cache ops
+		syn=N	Number of in-progress sync_cache() cache ops
+		atc=N	Number of in-progress attr_changed() cache ops
+		rap=N	Number of in-progress read_or_alloc_page() cache ops
+		ras=N	Number of in-progress read_or_alloc_pages() cache ops
+		alp=N	Number of in-progress allocate_page() cache ops
+		als=N	Number of in-progress allocate_pages() cache ops
+		wrp=N	Number of in-progress write_page() cache ops
+		ucp=N	Number of in-progress uncache_page() cache ops
+		dsp=N	Number of in-progress dissociate_pages() cache ops
+	CacheEv	nsp=N	Number of object lookups/creations rejected due to lack of space
+		stl=N	Number of stale objects deleted
+		rtr=N	Number of objects retired when relinquished
+		cul=N	Number of objects culled
+
+
+ (*) /proc/fs/fscache/histogram
+
+	cat /proc/fs/fscache/histogram
+	JIFS  SECS  OBJ INST  OP RUNS   OBJ RUNS  RETRV DLY RETRIEVLS
+	===== ===== ========= ========= ========= ========= =========
+
+     This shows the breakdown of the number of times each amount of time
+     between 0 jiffies and HZ-1 jiffies a variety of tasks took to run.  The
+     columns are as follows:
+
+	COLUMN		TIME MEASUREMENT
+	=======		=======================================================
+	OBJ INST	Length of time to instantiate an object
+	OP RUNS		Length of time a call to process an operation took
+	OBJ RUNS	Length of time a call to process an object event took
+	RETRV DLY	Time between an requesting a read and lookup completing
+	RETRIEVLS	Time between beginning and end of a retrieval
+
+     Each row shows the number of events that took a particular range of times.
+     Each step is 1 jiffy in size.  The JIFS column indicates the particular
+     jiffy range covered, and the SECS field the equivalent number of seconds.
+
+
+===========
+OBJECT LIST
+===========
+
+If CONFIG_FSCACHE_OBJECT_LIST is enabled, the FS-Cache facility will maintain a
+list of all the objects currently allocated and allow them to be viewed
+through:
+
+	/proc/fs/fscache/objects
+
+This will look something like:
+
+	[root@andromeda ~]# head /proc/fs/fscache/objects
+	OBJECT   PARENT   STAT CHLDN OPS OOP IPR EX READS EM EV F S | NETFS_COOKIE_DEF TY FL NETFS_DATA       OBJECT_KEY, AUX_DATA
+	======== ======== ==== ===== === === === == ===== == == = = | ================ == == ================ ================
+	   17e4b        2 ACTV     0   0   0   0  0     0 7b  4 0 0 | NFS.fh           DT  0 ffff88001dd82820 010006017edcf8bbc93b43298fdfbe71e50b57b13a172c0117f38472, e567634700000000000000000000000063f2404a000000000000000000000000c9030000000000000000000063f2404a
+	   1693a        2 ACTV     0   0   0   0  0     0 7b  4 0 0 | NFS.fh           DT  0 ffff88002db23380 010006017edcf8bbc93b43298fdfbe71e50b57b1e0162c01a2df0ea6, 420ebc4a000000000000000000000000420ebc4a0000000000000000000000000e1801000000000000000000420ebc4a
+
+where the first set of columns before the '|' describe the object:
+
+	COLUMN	DESCRIPTION
+	=======	===============================================================
+	OBJECT	Object debugging ID (appears as OBJ%x in some debug messages)
+	PARENT	Debugging ID of parent object
+	STAT	Object state
+	CHLDN	Number of child objects of this object
+	OPS	Number of outstanding operations on this object
+	OOP	Number of outstanding child object management operations
+	IPR
+	EX	Number of outstanding exclusive operations
+	READS	Number of outstanding read operations
+	EM	Object's event mask
+	EV	Events raised on this object
+	F	Object flags
+	S	Object work item busy state mask (1:pending 2:running)
+
+and the second set of columns describe the object's cookie, if present:
+
+	COLUMN		DESCRIPTION
+	===============	=======================================================
+	NETFS_COOKIE_DEF Name of netfs cookie definition
+	TY		Cookie type (IX - index, DT - data, hex - special)
+	FL		Cookie flags
+	NETFS_DATA	Netfs private data stored in the cookie
+	OBJECT_KEY	Object key	} 1 column, with separating comma
+	AUX_DATA	Object aux data	} presence may be configured
+
+The data shown may be filtered by attaching the a key to an appropriate keyring
+before viewing the file.  Something like:
+
+		keyctl add user fscache:objlist <restrictions> @s
+
+where <restrictions> are a selection of the following letters:
+
+	K	Show hexdump of object key (don't show if not given)
+	A	Show hexdump of object aux data (don't show if not given)
+
+and the following paired letters:
+
+	C	Show objects that have a cookie
+	c	Show objects that don't have a cookie
+	B	Show objects that are busy
+	b	Show objects that aren't busy
+	W	Show objects that have pending writes
+	w	Show objects that don't have pending writes
+	R	Show objects that have outstanding reads
+	r	Show objects that don't have outstanding reads
+	S	Show objects that have work queued
+	s	Show objects that don't have work queued
+
+If neither side of a letter pair is given, then both are implied.  For example:
+
+	keyctl add user fscache:objlist KB @s
+
+shows objects that are busy, and lists their object keys, but does not dump
+their auxiliary data.  It also implies "CcWwRrSs", but as 'B' is given, 'b' is
+not implied.
+
+By default all objects and all fields will be shown.
+
+
+=========
+DEBUGGING
+=========
+
+If CONFIG_FSCACHE_DEBUG is enabled, the FS-Cache facility can have runtime
+debugging enabled by adjusting the value in:
+
+	/sys/module/fscache/parameters/debug
+
+This is a bitmask of debugging streams to enable:
+
+	BIT	VALUE	STREAM				POINT
+	=======	=======	===============================	=======================
+	0	1	Cache management		Function entry trace
+	1	2					Function exit trace
+	2	4					General
+	3	8	Cookie management		Function entry trace
+	4	16					Function exit trace
+	5	32					General
+	6	64	Page handling			Function entry trace
+	7	128					Function exit trace
+	8	256					General
+	9	512	Operation management		Function entry trace
+	10	1024					Function exit trace
+	11	2048					General
+
+The appropriate set of values should be OR'd together and the result written to
+the control file.  For example:
+
+	echo $((1|8|64)) >/sys/module/fscache/parameters/debug
+
+will turn on all function entry debugging.
diff --git a/Documentation/filesystems/caching/netfs-api.txt b/Documentation/filesystems/caching/netfs-api.txt
new file mode 100644
index 000000000..2a6f7399c
--- /dev/null
+++ b/Documentation/filesystems/caching/netfs-api.txt
@@ -0,0 +1,910 @@
+			===============================
+			FS-CACHE NETWORK FILESYSTEM API
+			===============================
+
+There's an API by which a network filesystem can make use of the FS-Cache
+facilities.  This is based around a number of principles:
+
+ (1) Caches can store a number of different object types.  There are two main
+     object types: indices and files.  The first is a special type used by
+     FS-Cache to make finding objects faster and to make retiring of groups of
+     objects easier.
+
+ (2) Every index, file or other object is represented by a cookie.  This cookie
+     may or may not have anything associated with it, but the netfs doesn't
+     need to care.
+
+ (3) Barring the top-level index (one entry per cached netfs), the index
+     hierarchy for each netfs is structured according the whim of the netfs.
+
+This API is declared in <linux/fscache.h>.
+
+This document contains the following sections:
+
+	 (1) Network filesystem definition
+	 (2) Index definition
+	 (3) Object definition
+	 (4) Network filesystem (un)registration
+	 (5) Cache tag lookup
+	 (6) Index registration
+	 (7) Data file registration
+	 (8) Miscellaneous object registration
+ 	 (9) Setting the data file size
+	(10) Page alloc/read/write
+	(11) Page uncaching
+	(12) Index and data file consistency
+	(13) Cookie enablement
+	(14) Miscellaneous cookie operations
+	(15) Cookie unregistration
+	(16) Index invalidation
+	(17) Data file invalidation
+	(18) FS-Cache specific page flags.
+
+
+=============================
+NETWORK FILESYSTEM DEFINITION
+=============================
+
+FS-Cache needs a description of the network filesystem.  This is specified
+using a record of the following structure:
+
+	struct fscache_netfs {
+		uint32_t			version;
+		const char			*name;
+		struct fscache_cookie		*primary_index;
+		...
+	};
+
+This first two fields should be filled in before registration, and the third
+will be filled in by the registration function; any other fields should just be
+ignored and are for internal use only.
+
+The fields are:
+
+ (1) The name of the netfs (used as the key in the toplevel index).
+
+ (2) The version of the netfs (if the name matches but the version doesn't, the
+     entire in-cache hierarchy for this netfs will be scrapped and begun
+     afresh).
+
+ (3) The cookie representing the primary index will be allocated according to
+     another parameter passed into the registration function.
+
+For example, kAFS (linux/fs/afs/) uses the following definitions to describe
+itself:
+
+	struct fscache_netfs afs_cache_netfs = {
+		.version	= 0,
+		.name		= "afs",
+	};
+
+
+================
+INDEX DEFINITION
+================
+
+Indices are used for two purposes:
+
+ (1) To aid the finding of a file based on a series of keys (such as AFS's
+     "cell", "volume ID", "vnode ID").
+
+ (2) To make it easier to discard a subset of all the files cached based around
+     a particular key - for instance to mirror the removal of an AFS volume.
+
+However, since it's unlikely that any two netfs's are going to want to define
+their index hierarchies in quite the same way, FS-Cache tries to impose as few
+restraints as possible on how an index is structured and where it is placed in
+the tree.  The netfs can even mix indices and data files at the same level, but
+it's not recommended.
+
+Each index entry consists of a key of indeterminate length plus some auxiliary
+data, also of indeterminate length.
+
+There are some limits on indices:
+
+ (1) Any index containing non-index objects should be restricted to a single
+     cache.  Any such objects created within an index will be created in the
+     first cache only.  The cache in which an index is created can be
+     controlled by cache tags (see below).
+
+ (2) The entry data must be atomically journallable, so it is limited to about
+     400 bytes at present.  At least 400 bytes will be available.
+
+ (3) The depth of the index tree should be judged with care as the search
+     function is recursive.  Too many layers will run the kernel out of stack.
+
+
+=================
+OBJECT DEFINITION
+=================
+
+To define an object, a structure of the following type should be filled out:
+
+	struct fscache_cookie_def
+	{
+		uint8_t name[16];
+		uint8_t type;
+
+		struct fscache_cache_tag *(*select_cache)(
+			const void *parent_netfs_data,
+			const void *cookie_netfs_data);
+
+		enum fscache_checkaux (*check_aux)(void *cookie_netfs_data,
+						   const void *data,
+						   uint16_t datalen,
+						   loff_t object_size);
+
+		void (*get_context)(void *cookie_netfs_data, void *context);
+
+		void (*put_context)(void *cookie_netfs_data, void *context);
+
+		void (*mark_pages_cached)(void *cookie_netfs_data,
+					  struct address_space *mapping,
+					  struct pagevec *cached_pvec);
+	};
+
+This has the following fields:
+
+ (1) The type of the object [mandatory].
+
+     This is one of the following values:
+
+	(*) FSCACHE_COOKIE_TYPE_INDEX
+
+	    This defines an index, which is a special FS-Cache type.
+
+	(*) FSCACHE_COOKIE_TYPE_DATAFILE
+
+	    This defines an ordinary data file.
+
+	(*) Any other value between 2 and 255
+
+	    This defines an extraordinary object such as an XATTR.
+
+ (2) The name of the object type (NUL terminated unless all 16 chars are used)
+     [optional].
+
+ (3) A function to select the cache in which to store an index [optional].
+
+     This function is invoked when an index needs to be instantiated in a cache
+     during the instantiation of a non-index object.  Only the immediate index
+     parent for the non-index object will be queried.  Any indices above that
+     in the hierarchy may be stored in multiple caches.  This function does not
+     need to be supplied for any non-index object or any index that will only
+     have index children.
+
+     If this function is not supplied or if it returns NULL then the first
+     cache in the parent's list will be chosen, or failing that, the first
+     cache in the master list.
+
+ (4) A function to check the auxiliary data [optional].
+
+     This function will be called to check that a match found in the cache for
+     this object is valid.  For instance with AFS it could check the auxiliary
+     data against the data version number returned by the server to determine
+     whether the index entry in a cache is still valid.
+
+     If this function is absent, it will be assumed that matching objects in a
+     cache are always valid.
+
+     The function is also passed the cache's idea of the object size and may
+     use this to manage coherency also.
+
+     If present, the function should return one of the following values:
+
+	(*) FSCACHE_CHECKAUX_OKAY		- the entry is okay as is
+	(*) FSCACHE_CHECKAUX_NEEDS_UPDATE	- the entry requires update
+	(*) FSCACHE_CHECKAUX_OBSOLETE		- the entry should be deleted
+
+     This function can also be used to extract data from the auxiliary data in
+     the cache and copy it into the netfs's structures.
+
+ (5) A pair of functions to manage contexts for the completion callback
+     [optional].
+
+     The cache read/write functions are passed a context which is then passed
+     to the I/O completion callback function.  To ensure this context remains
+     valid until after the I/O completion is called, two functions may be
+     provided: one to get an extra reference on the context, and one to drop a
+     reference to it.
+
+     If the context is not used or is a type of object that won't go out of
+     scope, then these functions are not required.  These functions are not
+     required for indices as indices may not contain data.  These functions may
+     be called in interrupt context and so may not sleep.
+
+ (6) A function to mark a page as retaining cache metadata [optional].
+
+     This is called by the cache to indicate that it is retaining in-memory
+     information for this page and that the netfs should uncache the page when
+     it has finished.  This does not indicate whether there's data on the disk
+     or not.  Note that several pages at once may be presented for marking.
+
+     The PG_fscache bit is set on the pages before this function would be
+     called, so the function need not be provided if this is sufficient.
+
+     This function is not required for indices as they're not permitted data.
+
+ (7) A function to unmark all the pages retaining cache metadata [mandatory].
+
+     This is called by FS-Cache to indicate that a backing store is being
+     unbound from a cookie and that all the marks on the pages should be
+     cleared to prevent confusion.  Note that the cache will have torn down all
+     its tracking information so that the pages don't need to be explicitly
+     uncached.
+
+     This function is not required for indices as they're not permitted data.
+
+
+===================================
+NETWORK FILESYSTEM (UN)REGISTRATION
+===================================
+
+The first step is to declare the network filesystem to the cache.  This also
+involves specifying the layout of the primary index (for AFS, this would be the
+"cell" level).
+
+The registration function is:
+
+	int fscache_register_netfs(struct fscache_netfs *netfs);
+
+It just takes a pointer to the netfs definition.  It returns 0 or an error as
+appropriate.
+
+For kAFS, registration is done as follows:
+
+	ret = fscache_register_netfs(&afs_cache_netfs);
+
+The last step is, of course, unregistration:
+
+	void fscache_unregister_netfs(struct fscache_netfs *netfs);
+
+
+================
+CACHE TAG LOOKUP
+================
+
+FS-Cache permits the use of more than one cache.  To permit particular index
+subtrees to be bound to particular caches, the second step is to look up cache
+representation tags.  This step is optional; it can be left entirely up to
+FS-Cache as to which cache should be used.  The problem with doing that is that
+FS-Cache will always pick the first cache that was registered.
+
+To get the representation for a named tag:
+
+	struct fscache_cache_tag *fscache_lookup_cache_tag(const char *name);
+
+This takes a text string as the name and returns a representation of a tag.  It
+will never return an error.  It may return a dummy tag, however, if it runs out
+of memory; this will inhibit caching with this tag.
+
+Any representation so obtained must be released by passing it to this function:
+
+	void fscache_release_cache_tag(struct fscache_cache_tag *tag);
+
+The tag will be retrieved by FS-Cache when it calls the object definition
+operation select_cache().
+
+
+==================
+INDEX REGISTRATION
+==================
+
+The third step is to inform FS-Cache about part of an index hierarchy that can
+be used to locate files.  This is done by requesting a cookie for each index in
+the path to the file:
+
+	struct fscache_cookie *
+	fscache_acquire_cookie(struct fscache_cookie *parent,
+			       const struct fscache_object_def *def,
+			       const void *index_key,
+			       size_t index_key_len,
+			       const void *aux_data,
+			       size_t aux_data_len,
+			       void *netfs_data,
+			       loff_t object_size,
+			       bool enable);
+
+This function creates an index entry in the index represented by parent,
+filling in the index entry by calling the operations pointed to by def.
+
+A unique key that represents the object within the parent must be pointed to by
+index_key and is of length index_key_len.
+
+An optional blob of auxiliary data that is to be stored within the cache can be
+pointed to with aux_data and should be of length aux_data_len.  This would
+typically be used for storing coherency data.
+
+The netfs may pass an arbitrary value in netfs_data and this will be presented
+to it in the event of any calling back.  This may also be used in tracing or
+logging of messages.
+
+The cache tracks the size of the data attached to an object and this set to be
+object_size.  For indices, this should be 0.  This value will be passed to the
+->check_aux() callback.
+
+Note that this function never returns an error - all errors are handled
+internally.  It may, however, return NULL to indicate no cookie.  It is quite
+acceptable to pass this token back to this function as the parent to another
+acquisition (or even to the relinquish cookie, read page and write page
+functions - see below).
+
+Note also that no indices are actually created in a cache until a non-index
+object needs to be created somewhere down the hierarchy.  Furthermore, an index
+may be created in several different caches independently at different times.
+This is all handled transparently, and the netfs doesn't see any of it.
+
+A cookie will be created in the disabled state if enabled is false.  A cookie
+must be enabled to do anything with it.  A disabled cookie can be enabled by
+calling fscache_enable_cookie() (see below).
+
+For example, with AFS, a cell would be added to the primary index.  This index
+entry would have a dependent inode containing volume mappings within this cell:
+
+	cell->cache =
+		fscache_acquire_cookie(afs_cache_netfs.primary_index,
+				       &afs_cell_cache_index_def,
+				       cell->name, strlen(cell->name),
+				       NULL, 0,
+				       cell, 0, true);
+
+And then a particular volume could be added to that index by ID, creating
+another index for vnodes (AFS inode equivalents):
+
+	volume->cache =
+		fscache_acquire_cookie(volume->cell->cache,
+				       &afs_volume_cache_index_def,
+				       &volume->vid, sizeof(volume->vid),
+				       NULL, 0,
+				       volume, 0, true);
+
+
+======================
+DATA FILE REGISTRATION
+======================
+
+The fourth step is to request a data file be created in the cache.  This is
+identical to index cookie acquisition.  The only difference is that the type in
+the object definition should be something other than index type.
+
+	vnode->cache =
+		fscache_acquire_cookie(volume->cache,
+				       &afs_vnode_cache_object_def,
+				       &key, sizeof(key),
+				       &aux, sizeof(aux),
+				       vnode, vnode->status.size, true);
+
+
+=================================
+MISCELLANEOUS OBJECT REGISTRATION
+=================================
+
+An optional step is to request an object of miscellaneous type be created in
+the cache.  This is almost identical to index cookie acquisition.  The only
+difference is that the type in the object definition should be something other
+than index type.  Whilst the parent object could be an index, it's more likely
+it would be some other type of object such as a data file.
+
+	xattr->cache =
+		fscache_acquire_cookie(vnode->cache,
+				       &afs_xattr_cache_object_def,
+				       &xattr->name, strlen(xattr->name),
+				       NULL, 0,
+				       xattr, strlen(xattr->val), true);
+
+Miscellaneous objects might be used to store extended attributes or directory
+entries for example.
+
+
+==========================
+SETTING THE DATA FILE SIZE
+==========================
+
+The fifth step is to set the physical attributes of the file, such as its size.
+This doesn't automatically reserve any space in the cache, but permits the
+cache to adjust its metadata for data tracking appropriately:
+
+	int fscache_attr_changed(struct fscache_cookie *cookie);
+
+The cache will return -ENOBUFS if there is no backing cache or if there is no
+space to allocate any extra metadata required in the cache.
+
+Note that attempts to read or write data pages in the cache over this size may
+be rebuffed with -ENOBUFS.
+
+This operation schedules an attribute adjustment to happen asynchronously at
+some point in the future, and as such, it may happen after the function returns
+to the caller.  The attribute adjustment excludes read and write operations.
+
+
+=====================
+PAGE ALLOC/READ/WRITE
+=====================
+
+And the sixth step is to store and retrieve pages in the cache.  There are
+three functions that are used to do this.
+
+Note:
+
+ (1) A page should not be re-read or re-allocated without uncaching it first.
+
+ (2) A read or allocated page must be uncached when the netfs page is released
+     from the pagecache.
+
+ (3) A page should only be written to the cache if previous read or allocated.
+
+This permits the cache to maintain its page tracking in proper order.
+
+
+PAGE READ
+---------
+
+Firstly, the netfs should ask FS-Cache to examine the caches and read the
+contents cached for a particular page of a particular file if present, or else
+allocate space to store the contents if not:
+
+	typedef
+	void (*fscache_rw_complete_t)(struct page *page,
+				      void *context,
+				      int error);
+
+	int fscache_read_or_alloc_page(struct fscache_cookie *cookie,
+				       struct page *page,
+				       fscache_rw_complete_t end_io_func,
+				       void *context,
+				       gfp_t gfp);
+
+The cookie argument must specify a cookie for an object that isn't an index,
+the page specified will have the data loaded into it (and is also used to
+specify the page number), and the gfp argument is used to control how any
+memory allocations made are satisfied.
+
+If the cookie indicates the inode is not cached:
+
+ (1) The function will return -ENOBUFS.
+
+Else if there's a copy of the page resident in the cache:
+
+ (1) The mark_pages_cached() cookie operation will be called on that page.
+
+ (2) The function will submit a request to read the data from the cache's
+     backing device directly into the page specified.
+
+ (3) The function will return 0.
+
+ (4) When the read is complete, end_io_func() will be invoked with:
+
+     (*) The netfs data supplied when the cookie was created.
+
+     (*) The page descriptor.
+
+     (*) The context argument passed to the above function.  This will be
+         maintained with the get_context/put_context functions mentioned above.
+
+     (*) An argument that's 0 on success or negative for an error code.
+
+     If an error occurs, it should be assumed that the page contains no usable
+     data.  fscache_readpages_cancel() may need to be called.
+
+     end_io_func() will be called in process context if the read is results in
+     an error, but it might be called in interrupt context if the read is
+     successful.
+
+Otherwise, if there's not a copy available in cache, but the cache may be able
+to store the page:
+
+ (1) The mark_pages_cached() cookie operation will be called on that page.
+
+ (2) A block may be reserved in the cache and attached to the object at the
+     appropriate place.
+
+ (3) The function will return -ENODATA.
+
+This function may also return -ENOMEM or -EINTR, in which case it won't have
+read any data from the cache.
+
+
+PAGE ALLOCATE
+-------------
+
+Alternatively, if there's not expected to be any data in the cache for a page
+because the file has been extended, a block can simply be allocated instead:
+
+	int fscache_alloc_page(struct fscache_cookie *cookie,
+			       struct page *page,
+			       gfp_t gfp);
+
+This is similar to the fscache_read_or_alloc_page() function, except that it
+never reads from the cache.  It will return 0 if a block has been allocated,
+rather than -ENODATA as the other would.  One or the other must be performed
+before writing to the cache.
+
+The mark_pages_cached() cookie operation will be called on the page if
+successful.
+
+
+PAGE WRITE
+----------
+
+Secondly, if the netfs changes the contents of the page (either due to an
+initial download or if a user performs a write), then the page should be
+written back to the cache:
+
+	int fscache_write_page(struct fscache_cookie *cookie,
+			       struct page *page,
+			       loff_t object_size,
+			       gfp_t gfp);
+
+The cookie argument must specify a data file cookie, the page specified should
+contain the data to be written (and is also used to specify the page number),
+object_size is the revised size of the object and the gfp argument is used to
+control how any memory allocations made are satisfied.
+
+The page must have first been read or allocated successfully and must not have
+been uncached before writing is performed.
+
+If the cookie indicates the inode is not cached then:
+
+ (1) The function will return -ENOBUFS.
+
+Else if space can be allocated in the cache to hold this page:
+
+ (1) PG_fscache_write will be set on the page.
+
+ (2) The function will submit a request to write the data to cache's backing
+     device directly from the page specified.
+
+ (3) The function will return 0.
+
+ (4) When the write is complete PG_fscache_write is cleared on the page and
+     anyone waiting for that bit will be woken up.
+
+Else if there's no space available in the cache, -ENOBUFS will be returned.  It
+is also possible for the PG_fscache_write bit to be cleared when no write took
+place if unforeseen circumstances arose (such as a disk error).
+
+Writing takes place asynchronously.
+
+
+MULTIPLE PAGE READ
+------------------
+
+A facility is provided to read several pages at once, as requested by the
+readpages() address space operation:
+
+	int fscache_read_or_alloc_pages(struct fscache_cookie *cookie,
+					struct address_space *mapping,
+					struct list_head *pages,
+					int *nr_pages,
+					fscache_rw_complete_t end_io_func,
+					void *context,
+					gfp_t gfp);
+
+This works in a similar way to fscache_read_or_alloc_page(), except:
+
+ (1) Any page it can retrieve data for is removed from pages and nr_pages and
+     dispatched for reading to the disk.  Reads of adjacent pages on disk may
+     be merged for greater efficiency.
+
+ (2) The mark_pages_cached() cookie operation will be called on several pages
+     at once if they're being read or allocated.
+
+ (3) If there was an general error, then that error will be returned.
+
+     Else if some pages couldn't be allocated or read, then -ENOBUFS will be
+     returned.
+
+     Else if some pages couldn't be read but were allocated, then -ENODATA will
+     be returned.
+
+     Otherwise, if all pages had reads dispatched, then 0 will be returned, the
+     list will be empty and *nr_pages will be 0.
+
+ (4) end_io_func will be called once for each page being read as the reads
+     complete.  It will be called in process context if error != 0, but it may
+     be called in interrupt context if there is no error.
+
+Note that a return of -ENODATA, -ENOBUFS or any other error does not preclude
+some of the pages being read and some being allocated.  Those pages will have
+been marked appropriately and will need uncaching.
+
+
+CANCELLATION OF UNREAD PAGES
+----------------------------
+
+If one or more pages are passed to fscache_read_or_alloc_pages() but not then
+read from the cache and also not read from the underlying filesystem then
+those pages will need to have any marks and reservations removed.  This can be
+done by calling:
+
+	void fscache_readpages_cancel(struct fscache_cookie *cookie,
+				      struct list_head *pages);
+
+prior to returning to the caller.  The cookie argument should be as passed to
+fscache_read_or_alloc_pages().  Every page in the pages list will be examined
+and any that have PG_fscache set will be uncached.
+
+
+==============
+PAGE UNCACHING
+==============
+
+To uncache a page, this function should be called:
+
+	void fscache_uncache_page(struct fscache_cookie *cookie,
+				  struct page *page);
+
+This function permits the cache to release any in-memory representation it
+might be holding for this netfs page.  This function must be called once for
+each page on which the read or write page functions above have been called to
+make sure the cache's in-memory tracking information gets torn down.
+
+Note that pages can't be explicitly deleted from the a data file.  The whole
+data file must be retired (see the relinquish cookie function below).
+
+Furthermore, note that this does not cancel the asynchronous read or write
+operation started by the read/alloc and write functions, so the page
+invalidation functions must use:
+
+	bool fscache_check_page_write(struct fscache_cookie *cookie,
+				      struct page *page);
+
+to see if a page is being written to the cache, and:
+
+	void fscache_wait_on_page_write(struct fscache_cookie *cookie,
+					struct page *page);
+
+to wait for it to finish if it is.
+
+
+When releasepage() is being implemented, a special FS-Cache function exists to
+manage the heuristics of coping with vmscan trying to eject pages, which may
+conflict with the cache trying to write pages to the cache (which may itself
+need to allocate memory):
+
+	bool fscache_maybe_release_page(struct fscache_cookie *cookie,
+					struct page *page,
+					gfp_t gfp);
+
+This takes the netfs cookie, and the page and gfp arguments as supplied to
+releasepage().  It will return false if the page cannot be released yet for
+some reason and if it returns true, the page has been uncached and can now be
+released.
+
+To make a page available for release, this function may wait for an outstanding
+storage request to complete, or it may attempt to cancel the storage request -
+in which case the page will not be stored in the cache this time.
+
+
+BULK INODE PAGE UNCACHE
+-----------------------
+
+A convenience routine is provided to perform an uncache on all the pages
+attached to an inode.  This assumes that the pages on the inode correspond on a
+1:1 basis with the pages in the cache.
+
+	void fscache_uncache_all_inode_pages(struct fscache_cookie *cookie,
+					     struct inode *inode);
+
+This takes the netfs cookie that the pages were cached with and the inode that
+the pages are attached to.  This function will wait for pages to finish being
+written to the cache and for the cache to finish with the page generally.  No
+error is returned.
+
+
+===============================
+INDEX AND DATA FILE CONSISTENCY
+===============================
+
+To find out whether auxiliary data for an object is up to data within the
+cache, the following function can be called:
+
+	int fscache_check_consistency(struct fscache_cookie *cookie,
+				      const void *aux_data);
+
+This will call back to the netfs to check whether the auxiliary data associated
+with a cookie is correct; if aux_data is non-NULL, it will update the auxiliary
+data buffer first.  It returns 0 if it is and -ESTALE if it isn't; it may also
+return -ENOMEM and -ERESTARTSYS.
+
+To request an update of the index data for an index or other object, the
+following function should be called:
+
+	void fscache_update_cookie(struct fscache_cookie *cookie,
+				   const void *aux_data);
+
+This function will update the cookie's auxiliary data buffer from aux_data if
+that is non-NULL and then schedule this to be stored on disk.  The update
+method in the parent index definition will be called to transfer the data.
+
+Note that partial updates may happen automatically at other times, such as when
+data blocks are added to a data file object.
+
+
+=================
+COOKIE ENABLEMENT
+=================
+
+Cookies exist in one of two states: enabled and disabled.  If a cookie is
+disabled, it ignores all attempts to acquire child cookies; check, update or
+invalidate its state; allocate, read or write backing pages - though it is
+still possible to uncache pages and relinquish the cookie.
+
+The initial enablement state is set by fscache_acquire_cookie(), but the cookie
+can be enabled or disabled later.  To disable a cookie, call:
+
+	void fscache_disable_cookie(struct fscache_cookie *cookie,
+				    const void *aux_data,
+    				    bool invalidate);
+
+If the cookie is not already disabled, this locks the cookie against other
+enable and disable ops, marks the cookie as being disabled, discards or
+invalidates any backing objects and waits for cessation of activity on any
+associated object before unlocking the cookie.
+
+All possible failures are handled internally.  The caller should consider
+calling fscache_uncache_all_inode_pages() afterwards to make sure all page
+markings are cleared up.
+
+Cookies can be enabled or reenabled with:
+
+    	void fscache_enable_cookie(struct fscache_cookie *cookie,
+				   const void *aux_data,
+				   loff_t object_size,
+    				   bool (*can_enable)(void *data),
+    				   void *data)
+
+If the cookie is not already enabled, this locks the cookie against other
+enable and disable ops, invokes can_enable() and, if the cookie is not an index
+cookie, will begin the procedure of acquiring backing objects.
+
+The optional can_enable() function is passed the data argument and returns a
+ruling as to whether or not enablement should actually be permitted to begin.
+
+All possible failures are handled internally.  The cookie will only be marked
+as enabled if provisional backing objects are allocated.
+
+The object's data size is updated from object_size and is passed to the
+->check_aux() function.
+
+In both cases, the cookie's auxiliary data buffer is updated from aux_data if
+that is non-NULL inside the enablement lock before proceeding.
+
+
+===============================
+MISCELLANEOUS COOKIE OPERATIONS
+===============================
+
+There are a number of operations that can be used to control cookies:
+
+ (*) Cookie pinning:
+
+	int fscache_pin_cookie(struct fscache_cookie *cookie);
+	void fscache_unpin_cookie(struct fscache_cookie *cookie);
+
+     These operations permit data cookies to be pinned into the cache and to
+     have the pinning removed.  They are not permitted on index cookies.
+
+     The pinning function will return 0 if successful, -ENOBUFS in the cookie
+     isn't backed by a cache, -EOPNOTSUPP if the cache doesn't support pinning,
+     -ENOSPC if there isn't enough space to honour the operation, -ENOMEM or
+     -EIO if there's any other problem.
+
+ (*) Data space reservation:
+
+	int fscache_reserve_space(struct fscache_cookie *cookie, loff_t size);
+
+     This permits a netfs to request cache space be reserved to store up to the
+     given amount of a file.  It is permitted to ask for more than the current
+     size of the file to allow for future file expansion.
+
+     If size is given as zero then the reservation will be cancelled.
+
+     The function will return 0 if successful, -ENOBUFS in the cookie isn't
+     backed by a cache, -EOPNOTSUPP if the cache doesn't support reservations,
+     -ENOSPC if there isn't enough space to honour the operation, -ENOMEM or
+     -EIO if there's any other problem.
+
+     Note that this doesn't pin an object in a cache; it can still be culled to
+     make space if it's not in use.
+
+
+=====================
+COOKIE UNREGISTRATION
+=====================
+
+To get rid of a cookie, this function should be called.
+
+	void fscache_relinquish_cookie(struct fscache_cookie *cookie,
+				       const void *aux_data,
+				       bool retire);
+
+If retire is non-zero, then the object will be marked for recycling, and all
+copies of it will be removed from all active caches in which it is present.
+Not only that but all child objects will also be retired.
+
+If retire is zero, then the object may be available again when next the
+acquisition function is called.  Retirement here will overrule the pinning on a
+cookie.
+
+The cookie's auxiliary data will be updated from aux_data if that is non-NULL
+so that the cache can lazily update it on disk.
+
+One very important note - relinquish must NOT be called for a cookie unless all
+the cookies for "child" indices, objects and pages have been relinquished
+first.
+
+
+==================
+INDEX INVALIDATION
+==================
+
+There is no direct way to invalidate an index subtree.  To do this, the caller
+should relinquish and retire the cookie they have, and then acquire a new one.
+
+
+======================
+DATA FILE INVALIDATION
+======================
+
+Sometimes it will be necessary to invalidate an object that contains data.
+Typically this will be necessary when the server tells the netfs of a foreign
+change - at which point the netfs has to throw away all the state it had for an
+inode and reload from the server.
+
+To indicate that a cache object should be invalidated, the following function
+can be called:
+
+	void fscache_invalidate(struct fscache_cookie *cookie);
+
+This can be called with spinlocks held as it defers the work to a thread pool.
+All extant storage, retrieval and attribute change ops at this point are
+cancelled and discarded.  Some future operations will be rejected until the
+cache has had a chance to insert a barrier in the operations queue.  After
+that, operations will be queued again behind the invalidation operation.
+
+The invalidation operation will perform an attribute change operation and an
+auxiliary data update operation as it is very likely these will have changed.
+
+Using the following function, the netfs can wait for the invalidation operation
+to have reached a point at which it can start submitting ordinary operations
+once again:
+
+	void fscache_wait_on_invalidate(struct fscache_cookie *cookie);
+
+
+===========================
+FS-CACHE SPECIFIC PAGE FLAG
+===========================
+
+FS-Cache makes use of a page flag, PG_private_2, for its own purpose.  This is
+given the alternative name PG_fscache.
+
+PG_fscache is used to indicate that the page is known by the cache, and that
+the cache must be informed if the page is going to go away.  It's an indication
+to the netfs that the cache has an interest in this page, where an interest may
+be a pointer to it, resources allocated or reserved for it, or I/O in progress
+upon it.
+
+The netfs can use this information in methods such as releasepage() to
+determine whether it needs to uncache a page or update it.
+
+Furthermore, if this bit is set, releasepage() and invalidatepage() operations
+will be called on a page to get rid of it, even if PG_private is not set.  This
+allows caching to attempted on a page before read_cache_pages() to be called
+after fscache_read_or_alloc_pages() as the former will try and release pages it
+was given under certain circumstances.
+
+This bit does not overlap with such as PG_private.  This means that FS-Cache
+can be used with a filesystem that uses the block buffering code.
+
+There are a number of operations defined on this flag:
+
+	int PageFsCache(struct page *page);
+	void SetPageFsCache(struct page *page)
+	void ClearPageFsCache(struct page *page)
+	int TestSetPageFsCache(struct page *page)
+	int TestClearPageFsCache(struct page *page)
+
+These functions are bit test, bit set, bit clear, bit test and set and bit
+test and clear operations on PG_fscache.
diff --git a/Documentation/filesystems/caching/object.txt b/Documentation/filesystems/caching/object.txt
new file mode 100644
index 000000000..100ff4112
--- /dev/null
+++ b/Documentation/filesystems/caching/object.txt
@@ -0,0 +1,320 @@
+	     ====================================================
+	     IN-KERNEL CACHE OBJECT REPRESENTATION AND MANAGEMENT
+	     ====================================================
+
+By: David Howells <dhowells@redhat.com>
+
+Contents:
+
+ (*) Representation
+
+ (*) Object management state machine.
+
+     - Provision of cpu time.
+     - Locking simplification.
+
+ (*) The set of states.
+
+ (*) The set of events.
+
+
+==============
+REPRESENTATION
+==============
+
+FS-Cache maintains an in-kernel representation of each object that a netfs is
+currently interested in.  Such objects are represented by the fscache_cookie
+struct and are referred to as cookies.
+
+FS-Cache also maintains a separate in-kernel representation of the objects that
+a cache backend is currently actively caching.  Such objects are represented by
+the fscache_object struct.  The cache backends allocate these upon request, and
+are expected to embed them in their own representations.  These are referred to
+as objects.
+
+There is a 1:N relationship between cookies and objects.  A cookie may be
+represented by multiple objects - an index may exist in more than one cache -
+or even by no objects (it may not be cached).
+
+Furthermore, both cookies and objects are hierarchical.  The two hierarchies
+correspond, but the cookies tree is a superset of the union of the object trees
+of multiple caches:
+
+	    NETFS INDEX TREE               :      CACHE 1     :      CACHE 2
+	                                   :                  :
+	                                   :   +-----------+  :
+	                          +----------->|  IObject  |  :
+	      +-----------+       |        :   +-----------+  :
+	      |  ICookie  |-------+        :         |        :
+	      +-----------+       |        :         |        :   +-----------+
+	            |             +------------------------------>|  IObject  |
+	            |                      :         |        :   +-----------+
+	            |                      :         V        :         |
+	            |                      :   +-----------+  :         |
+	            V             +----------->|  IObject  |  :         |
+	      +-----------+       |        :   +-----------+  :         |
+	      |  ICookie  |-------+        :         |        :         V
+	      +-----------+       |        :         |        :   +-----------+
+	            |             +------------------------------>|  IObject  |
+	      +-----+-----+                :         |        :   +-----------+
+	      |           |                :         |        :         |
+	      V           |                :         V        :         |
+	+-----------+     |                :   +-----------+  :         |
+	|  ICookie  |------------------------->|  IObject  |  :         |
+	+-----------+     |                :   +-----------+  :         |
+	      |           V                :         |        :         V
+	      |     +-----------+          :         |        :   +-----------+
+	      |     |  ICookie  |-------------------------------->|  IObject  |
+	      |     +-----------+          :         |        :   +-----------+
+	      V           |                :         V        :         |
+	+-----------+     |                :   +-----------+  :         |
+	|  DCookie  |------------------------->|  DObject  |  :         |
+	+-----------+     |                :   +-----------+  :         |
+	                  |                :                  :         |
+	          +-------+-------+        :                  :         |
+	          |               |        :                  :         |
+	          V               V        :                  :         V
+	    +-----------+   +-----------+  :                  :   +-----------+
+	    |  DCookie  |   |  DCookie  |------------------------>|  DObject  |
+	    +-----------+   +-----------+  :                  :   +-----------+
+	                                   :                  :
+
+In the above illustration, ICookie and IObject represent indices and DCookie
+and DObject represent data storage objects.  Indices may have representation in
+multiple caches, but currently, non-index objects may not.  Objects of any type
+may also be entirely unrepresented.
+
+As far as the netfs API goes, the netfs is only actually permitted to see
+pointers to the cookies.  The cookies themselves and any objects attached to
+those cookies are hidden from it.
+
+
+===============================
+OBJECT MANAGEMENT STATE MACHINE
+===============================
+
+Within FS-Cache, each active object is managed by its own individual state
+machine.  The state for an object is kept in the fscache_object struct, in
+object->state.  A cookie may point to a set of objects that are in different
+states.
+
+Each state has an action associated with it that is invoked when the machine
+wakes up in that state.  There are four logical sets of states:
+
+ (1) Preparation: states that wait for the parent objects to become ready.  The
+     representations are hierarchical, and it is expected that an object must
+     be created or accessed with respect to its parent object.
+
+ (2) Initialisation: states that perform lookups in the cache and validate
+     what's found and that create on disk any missing metadata.
+
+ (3) Normal running: states that allow netfs operations on objects to proceed
+     and that update the state of objects.
+
+ (4) Termination: states that detach objects from their netfs cookies, that
+     delete objects from disk, that handle disk and system errors and that free
+     up in-memory resources.
+
+
+In most cases, transitioning between states is in response to signalled events.
+When a state has finished processing, it will usually set the mask of events in
+which it is interested (object->event_mask) and relinquish the worker thread.
+Then when an event is raised (by calling fscache_raise_event()), if the event
+is not masked, the object will be queued for processing (by calling
+fscache_enqueue_object()).
+
+
+PROVISION OF CPU TIME
+---------------------
+
+The work to be done by the various states was given CPU time by the threads of
+the slow work facility.  This was used in preference to the workqueue facility
+because:
+
+ (1) Threads may be completely occupied for very long periods of time by a
+     particular work item.  These state actions may be doing sequences of
+     synchronous, journalled disk accesses (lookup, mkdir, create, setxattr,
+     getxattr, truncate, unlink, rmdir, rename).
+
+ (2) Threads may do little actual work, but may rather spend a lot of time
+     sleeping on I/O.  This means that single-threaded and 1-per-CPU-threaded
+     workqueues don't necessarily have the right numbers of threads.
+
+
+LOCKING SIMPLIFICATION
+----------------------
+
+Because only one worker thread may be operating on any particular object's
+state machine at once, this simplifies the locking, particularly with respect
+to disconnecting the netfs's representation of a cache object (fscache_cookie)
+from the cache backend's representation (fscache_object) - which may be
+requested from either end.
+
+
+=================
+THE SET OF STATES
+=================
+
+The object state machine has a set of states that it can be in.  There are
+preparation states in which the object sets itself up and waits for its parent
+object to transit to a state that allows access to its children:
+
+ (1) State FSCACHE_OBJECT_INIT.
+
+     Initialise the object and wait for the parent object to become active.  In
+     the cache, it is expected that it will not be possible to look an object
+     up from the parent object, until that parent object itself has been looked
+     up.
+
+There are initialisation states in which the object sets itself up and accesses
+disk for the object metadata:
+
+ (2) State FSCACHE_OBJECT_LOOKING_UP.
+
+     Look up the object on disk, using the parent as a starting point.
+     FS-Cache expects the cache backend to probe the cache to see whether this
+     object is represented there, and if it is, to see if it's valid (coherency
+     management).
+
+     The cache should call fscache_object_lookup_negative() to indicate lookup
+     failure for whatever reason, and should call fscache_obtained_object() to
+     indicate success.
+
+     At the completion of lookup, FS-Cache will let the netfs go ahead with
+     read operations, no matter whether the file is yet cached.  If not yet
+     cached, read operations will be immediately rejected with ENODATA until
+     the first known page is uncached - as to that point there can be no data
+     to be read out of the cache for that file that isn't currently also held
+     in the pagecache.
+
+ (3) State FSCACHE_OBJECT_CREATING.
+
+     Create an object on disk, using the parent as a starting point.  This
+     happens if the lookup failed to find the object, or if the object's
+     coherency data indicated what's on disk is out of date.  In this state,
+     FS-Cache expects the cache to create
+
+     The cache should call fscache_obtained_object() if creation completes
+     successfully, fscache_object_lookup_negative() otherwise.
+
+     At the completion of creation, FS-Cache will start processing write
+     operations the netfs has queued for an object.  If creation failed, the
+     write ops will be transparently discarded, and nothing recorded in the
+     cache.
+
+There are some normal running states in which the object spends its time
+servicing netfs requests:
+
+ (4) State FSCACHE_OBJECT_AVAILABLE.
+
+     A transient state in which pending operations are started, child objects
+     are permitted to advance from FSCACHE_OBJECT_INIT state, and temporary
+     lookup data is freed.
+
+ (5) State FSCACHE_OBJECT_ACTIVE.
+
+     The normal running state.  In this state, requests the netfs makes will be
+     passed on to the cache.
+
+ (6) State FSCACHE_OBJECT_INVALIDATING.
+
+     The object is undergoing invalidation.  When the state comes here, it
+     discards all pending read, write and attribute change operations as it is
+     going to clear out the cache entirely and reinitialise it.  It will then
+     continue to the FSCACHE_OBJECT_UPDATING state.
+
+ (7) State FSCACHE_OBJECT_UPDATING.
+
+     The state machine comes here to update the object in the cache from the
+     netfs's records.  This involves updating the auxiliary data that is used
+     to maintain coherency.
+
+And there are terminal states in which an object cleans itself up, deallocates
+memory and potentially deletes stuff from disk:
+
+ (8) State FSCACHE_OBJECT_LC_DYING.
+
+     The object comes here if it is dying because of a lookup or creation
+     error.  This would be due to a disk error or system error of some sort.
+     Temporary data is cleaned up, and the parent is released.
+
+ (9) State FSCACHE_OBJECT_DYING.
+
+     The object comes here if it is dying due to an error, because its parent
+     cookie has been relinquished by the netfs or because the cache is being
+     withdrawn.
+
+     Any child objects waiting on this one are given CPU time so that they too
+     can destroy themselves.  This object waits for all its children to go away
+     before advancing to the next state.
+
+(10) State FSCACHE_OBJECT_ABORT_INIT.
+
+     The object comes to this state if it was waiting on its parent in
+     FSCACHE_OBJECT_INIT, but its parent died.  The object will destroy itself
+     so that the parent may proceed from the FSCACHE_OBJECT_DYING state.
+
+(11) State FSCACHE_OBJECT_RELEASING.
+(12) State FSCACHE_OBJECT_RECYCLING.
+
+     The object comes to one of these two states when dying once it is rid of
+     all its children, if it is dying because the netfs relinquished its
+     cookie.  In the first state, the cached data is expected to persist, and
+     in the second it will be deleted.
+
+(13) State FSCACHE_OBJECT_WITHDRAWING.
+
+     The object transits to this state if the cache decides it wants to
+     withdraw the object from service, perhaps to make space, but also due to
+     error or just because the whole cache is being withdrawn.
+
+(14) State FSCACHE_OBJECT_DEAD.
+
+     The object transits to this state when the in-memory object record is
+     ready to be deleted.  The object processor shouldn't ever see an object in
+     this state.
+
+
+THE SET OF EVENTS
+-----------------
+
+There are a number of events that can be raised to an object state machine:
+
+ (*) FSCACHE_OBJECT_EV_UPDATE
+
+     The netfs requested that an object be updated.  The state machine will ask
+     the cache backend to update the object, and the cache backend will ask the
+     netfs for details of the change through its cookie definition ops.
+
+ (*) FSCACHE_OBJECT_EV_CLEARED
+
+     This is signalled in two circumstances:
+
+     (a) when an object's last child object is dropped and
+
+     (b) when the last operation outstanding on an object is completed.
+
+     This is used to proceed from the dying state.
+
+ (*) FSCACHE_OBJECT_EV_ERROR
+
+     This is signalled when an I/O error occurs during the processing of some
+     object.
+
+ (*) FSCACHE_OBJECT_EV_RELEASE
+ (*) FSCACHE_OBJECT_EV_RETIRE
+
+     These are signalled when the netfs relinquishes a cookie it was using.
+     The event selected depends on whether the netfs asks for the backing
+     object to be retired (deleted) or retained.
+
+ (*) FSCACHE_OBJECT_EV_WITHDRAW
+
+     This is signalled when the cache backend wants to withdraw an object.
+     This means that the object will have to be detached from the netfs's
+     cookie.
+
+Because the withdrawing releasing/retiring events are all handled by the object
+state machine, it doesn't matter if there's a collision with both ends trying
+to sever the connection at the same time.  The state machine can just pick
+which one it wants to honour, and that effects the other.
diff --git a/Documentation/filesystems/caching/operations.txt b/Documentation/filesystems/caching/operations.txt
new file mode 100644
index 000000000..a1c052cbb
--- /dev/null
+++ b/Documentation/filesystems/caching/operations.txt
@@ -0,0 +1,213 @@
+		       ================================
+		       ASYNCHRONOUS OPERATIONS HANDLING
+		       ================================
+
+By: David Howells <dhowells@redhat.com>
+
+Contents:
+
+ (*) Overview.
+
+ (*) Operation record initialisation.
+
+ (*) Parameters.
+
+ (*) Procedure.
+
+ (*) Asynchronous callback.
+
+
+========
+OVERVIEW
+========
+
+FS-Cache has an asynchronous operations handling facility that it uses for its
+data storage and retrieval routines.  Its operations are represented by
+fscache_operation structs, though these are usually embedded into some other
+structure.
+
+This facility is available to and expected to be be used by the cache backends,
+and FS-Cache will create operations and pass them off to the appropriate cache
+backend for completion.
+
+To make use of this facility, <linux/fscache-cache.h> should be #included.
+
+
+===============================
+OPERATION RECORD INITIALISATION
+===============================
+
+An operation is recorded in an fscache_operation struct:
+
+	struct fscache_operation {
+		union {
+			struct work_struct fast_work;
+			struct slow_work slow_work;
+		};
+		unsigned long		flags;
+		fscache_operation_processor_t processor;
+		...
+	};
+
+Someone wanting to issue an operation should allocate something with this
+struct embedded in it.  They should initialise it by calling:
+
+	void fscache_operation_init(struct fscache_operation *op,
+				    fscache_operation_release_t release);
+
+with the operation to be initialised and the release function to use.
+
+The op->flags parameter should be set to indicate the CPU time provision and
+the exclusivity (see the Parameters section).
+
+The op->fast_work, op->slow_work and op->processor flags should be set as
+appropriate for the CPU time provision (see the Parameters section).
+
+FSCACHE_OP_WAITING may be set in op->flags prior to each submission of the
+operation and waited for afterwards.
+
+
+==========
+PARAMETERS
+==========
+
+There are a number of parameters that can be set in the operation record's flag
+parameter.  There are three options for the provision of CPU time in these
+operations:
+
+ (1) The operation may be done synchronously (FSCACHE_OP_MYTHREAD).  A thread
+     may decide it wants to handle an operation itself without deferring it to
+     another thread.
+
+     This is, for example, used in read operations for calling readpages() on
+     the backing filesystem in CacheFiles.  Although readpages() does an
+     asynchronous data fetch, the determination of whether pages exist is done
+     synchronously - and the netfs does not proceed until this has been
+     determined.
+
+     If this option is to be used, FSCACHE_OP_WAITING must be set in op->flags
+     before submitting the operation, and the operating thread must wait for it
+     to be cleared before proceeding:
+
+		wait_on_bit(&op->flags, FSCACHE_OP_WAITING,
+			    TASK_UNINTERRUPTIBLE);
+
+
+ (2) The operation may be fast asynchronous (FSCACHE_OP_FAST), in which case it
+     will be given to keventd to process.  Such an operation is not permitted
+     to sleep on I/O.
+
+     This is, for example, used by CacheFiles to copy data from a backing fs
+     page to a netfs page after the backing fs has read the page in.
+
+     If this option is used, op->fast_work and op->processor must be
+     initialised before submitting the operation:
+
+		INIT_WORK(&op->fast_work, do_some_work);
+
+
+ (3) The operation may be slow asynchronous (FSCACHE_OP_SLOW), in which case it
+     will be given to the slow work facility to process.  Such an operation is
+     permitted to sleep on I/O.
+
+     This is, for example, used by FS-Cache to handle background writes of
+     pages that have just been fetched from a remote server.
+
+     If this option is used, op->slow_work and op->processor must be
+     initialised before submitting the operation:
+
+		fscache_operation_init_slow(op, processor)
+
+
+Furthermore, operations may be one of two types:
+
+ (1) Exclusive (FSCACHE_OP_EXCLUSIVE).  Operations of this type may not run in
+     conjunction with any other operation on the object being operated upon.
+
+     An example of this is the attribute change operation, in which the file
+     being written to may need truncation.
+
+ (2) Shareable.  Operations of this type may be running simultaneously.  It's
+     up to the operation implementation to prevent interference between other
+     operations running at the same time.
+
+
+=========
+PROCEDURE
+=========
+
+Operations are used through the following procedure:
+
+ (1) The submitting thread must allocate the operation and initialise it
+     itself.  Normally this would be part of a more specific structure with the
+     generic op embedded within.
+
+ (2) The submitting thread must then submit the operation for processing using
+     one of the following two functions:
+
+	int fscache_submit_op(struct fscache_object *object,
+			      struct fscache_operation *op);
+
+	int fscache_submit_exclusive_op(struct fscache_object *object,
+					struct fscache_operation *op);
+
+     The first function should be used to submit non-exclusive ops and the
+     second to submit exclusive ones.  The caller must still set the
+     FSCACHE_OP_EXCLUSIVE flag.
+
+     If successful, both functions will assign the operation to the specified
+     object and return 0.  -ENOBUFS will be returned if the object specified is
+     permanently unavailable.
+
+     The operation manager will defer operations on an object that is still
+     undergoing lookup or creation.  The operation will also be deferred if an
+     operation of conflicting exclusivity is in progress on the object.
+
+     If the operation is asynchronous, the manager will retain a reference to
+     it, so the caller should put their reference to it by passing it to:
+
+	void fscache_put_operation(struct fscache_operation *op);
+
+ (3) If the submitting thread wants to do the work itself, and has marked the
+     operation with FSCACHE_OP_MYTHREAD, then it should monitor
+     FSCACHE_OP_WAITING as described above and check the state of the object if
+     necessary (the object might have died whilst the thread was waiting).
+
+     When it has finished doing its processing, it should call
+     fscache_op_complete() and fscache_put_operation() on it.
+
+ (4) The operation holds an effective lock upon the object, preventing other
+     exclusive ops conflicting until it is released.  The operation can be
+     enqueued for further immediate asynchronous processing by adjusting the
+     CPU time provisioning option if necessary, eg:
+
+	op->flags &= ~FSCACHE_OP_TYPE;
+	op->flags |= ~FSCACHE_OP_FAST;
+
+     and calling:
+
+	void fscache_enqueue_operation(struct fscache_operation *op)
+
+     This can be used to allow other things to have use of the worker thread
+     pools.
+
+
+=====================
+ASYNCHRONOUS CALLBACK
+=====================
+
+When used in asynchronous mode, the worker thread pool will invoke the
+processor method with a pointer to the operation.  This should then get at the
+container struct by using container_of():
+
+	static void fscache_write_op(struct fscache_operation *_op)
+	{
+		struct fscache_storage *op =
+			container_of(_op, struct fscache_storage, op);
+	...
+	}
+
+The caller holds a reference on the operation, and will invoke
+fscache_put_operation() when the processor function returns.  The processor
+function is at liberty to call fscache_enqueue_operation() or to take extra
+references.
diff --git a/Documentation/filesystems/ceph.txt b/Documentation/filesystems/ceph.txt
new file mode 100644
index 000000000..8bf62240e
--- /dev/null
+++ b/Documentation/filesystems/ceph.txt
@@ -0,0 +1,165 @@
+Ceph Distributed File System
+============================
+
+Ceph is a distributed network file system designed to provide good
+performance, reliability, and scalability.
+
+Basic features include:
+
+ * POSIX semantics
+ * Seamless scaling from 1 to many thousands of nodes
+ * High availability and reliability.  No single point of failure.
+ * N-way replication of data across storage nodes
+ * Fast recovery from node failures
+ * Automatic rebalancing of data on node addition/removal
+ * Easy deployment: most FS components are userspace daemons
+
+Also,
+ * Flexible snapshots (on any directory)
+ * Recursive accounting (nested files, directories, bytes)
+
+In contrast to cluster filesystems like GFS, OCFS2, and GPFS that rely
+on symmetric access by all clients to shared block devices, Ceph
+separates data and metadata management into independent server
+clusters, similar to Lustre.  Unlike Lustre, however, metadata and
+storage nodes run entirely as user space daemons.  Storage nodes
+utilize btrfs to store data objects, leveraging its advanced features
+(checksumming, metadata replication, etc.).  File data is striped
+across storage nodes in large chunks to distribute workload and
+facilitate high throughputs.  When storage nodes fail, data is
+re-replicated in a distributed fashion by the storage nodes themselves
+(with some minimal coordination from a cluster monitor), making the
+system extremely efficient and scalable.
+
+Metadata servers effectively form a large, consistent, distributed
+in-memory cache above the file namespace that is extremely scalable,
+dynamically redistributes metadata in response to workload changes,
+and can tolerate arbitrary (well, non-Byzantine) node failures.  The
+metadata server takes a somewhat unconventional approach to metadata
+storage to significantly improve performance for common workloads.  In
+particular, inodes with only a single link are embedded in
+directories, allowing entire directories of dentries and inodes to be
+loaded into its cache with a single I/O operation.  The contents of
+extremely large directories can be fragmented and managed by
+independent metadata servers, allowing scalable concurrent access.
+
+The system offers automatic data rebalancing/migration when scaling
+from a small cluster of just a few nodes to many hundreds, without
+requiring an administrator carve the data set into static volumes or
+go through the tedious process of migrating data between servers.
+When the file system approaches full, new nodes can be easily added
+and things will "just work."
+
+Ceph includes flexible snapshot mechanism that allows a user to create
+a snapshot on any subdirectory (and its nested contents) in the
+system.  Snapshot creation and deletion are as simple as 'mkdir
+.snap/foo' and 'rmdir .snap/foo'.
+
+Ceph also provides some recursive accounting on directories for nested
+files and bytes.  That is, a 'getfattr -d foo' on any directory in the
+system will reveal the total number of nested regular files and
+subdirectories, and a summation of all nested file sizes.  This makes
+the identification of large disk space consumers relatively quick, as
+no 'du' or similar recursive scan of the file system is required.
+
+Finally, Ceph also allows quotas to be set on any directory in the system.
+The quota can restrict the number of bytes or the number of files stored
+beneath that point in the directory hierarchy.  Quotas can be set using
+extended attributes 'ceph.quota.max_files' and 'ceph.quota.max_bytes', eg:
+
+ setfattr -n ceph.quota.max_bytes -v 100000000 /some/dir
+ getfattr -n ceph.quota.max_bytes /some/dir
+
+A limitation of the current quotas implementation is that it relies on the
+cooperation of the client mounting the file system to stop writers when a
+limit is reached.  A modified or adversarial client cannot be prevented
+from writing as much data as it needs.
+
+Mount Syntax
+============
+
+The basic mount syntax is:
+
+ # mount -t ceph monip[:port][,monip2[:port]...]:/[subdir] mnt
+
+You only need to specify a single monitor, as the client will get the
+full list when it connects.  (However, if the monitor you specify
+happens to be down, the mount won't succeed.)  The port can be left
+off if the monitor is using the default.  So if the monitor is at
+1.2.3.4,
+
+ # mount -t ceph 1.2.3.4:/ /mnt/ceph
+
+is sufficient.  If /sbin/mount.ceph is installed, a hostname can be
+used instead of an IP address.
+
+
+
+Mount Options
+=============
+
+  ip=A.B.C.D[:N]
+	Specify the IP and/or port the client should bind to locally.
+	There is normally not much reason to do this.  If the IP is not
+	specified, the client's IP address is determined by looking at the
+	address its connection to the monitor originates from.
+
+  wsize=X
+	Specify the maximum write size in bytes.  Default: 16 MB.
+
+  rsize=X
+	Specify the maximum read size in bytes.  Default: 16 MB.
+
+  rasize=X
+	Specify the maximum readahead size in bytes.  Default: 8 MB.
+
+  mount_timeout=X
+	Specify the timeout value for mount (in seconds), in the case
+	of a non-responsive Ceph file system.  The default is 30
+	seconds.
+
+  rbytes
+	When stat() is called on a directory, set st_size to 'rbytes',
+	the summation of file sizes over all files nested beneath that
+	directory.  This is the default.
+
+  norbytes
+	When stat() is called on a directory, set st_size to the
+	number of entries in that directory.
+
+  nocrc
+	Disable CRC32C calculation for data writes.  If set, the storage node
+	must rely on TCP's error correction to detect data corruption
+	in the data payload.
+
+  dcache
+        Use the dcache contents to perform negative lookups and
+        readdir when the client has the entire directory contents in
+        its cache.  (This does not change correctness; the client uses
+        cached metadata only when a lease or capability ensures it is
+        valid.)
+
+  nodcache
+        Do not use the dcache as above.  This avoids a significant amount of
+        complex code, sacrificing performance without affecting correctness,
+        and is useful for tracking down bugs.
+
+  noasyncreaddir
+	Do not use the dcache as above for readdir.
+
+  noquotadf
+        Report overall filesystem usage in statfs instead of using the root
+        directory quota.
+
+More Information
+================
+
+For more information on Ceph, see the home page at
+	http://ceph.newdream.net/
+
+The Linux kernel client source tree is available at
+	git://ceph.newdream.net/git/ceph-client.git
+	git://git.kernel.org/pub/scm/linux/kernel/git/sage/ceph-client.git
+
+and the source for the full system is at
+	git://ceph.newdream.net/git/ceph.git
diff --git a/Documentation/filesystems/cifs/AUTHORS b/Documentation/filesystems/cifs/AUTHORS
new file mode 100644
index 000000000..75865da2c
--- /dev/null
+++ b/Documentation/filesystems/cifs/AUTHORS
@@ -0,0 +1,63 @@
+Original Author
+===============
+Steve French (sfrench@samba.org)
+
+The author wishes to express his appreciation and thanks to:
+Andrew Tridgell (Samba team) for his early suggestions about smb/cifs VFS
+improvements. Thanks to IBM for allowing me time and test resources to pursue
+this project, to Jim McDonough from IBM (and the Samba Team) for his help, to
+the IBM Linux JFS team for explaining many esoteric Linux filesystem features.
+Jeremy Allison of the Samba team has done invaluable work in adding the server
+side of the original CIFS Unix extensions and reviewing and implementing
+portions of the newer CIFS POSIX extensions into the Samba 3 file server. Thank
+Dave Boutcher of IBM Rochester (author of the OS/400 smb/cifs filesystem client)
+for proving years ago that very good smb/cifs clients could be done on Unix-like
+operating systems.  Volker Lendecke, Andrew Tridgell, Urban Widmark, John 
+Newbigin and others for their work on the Linux smbfs module.  Thanks to
+the other members of the Storage Network Industry Association CIFS Technical
+Workgroup for their work specifying this highly complex protocol and finally
+thanks to the Samba team for their technical advice and encouragement.
+
+Patch Contributors
+------------------
+Zwane Mwaikambo
+Andi Kleen
+Amrut Joshi
+Shobhit Dayal
+Sergey Vlasov
+Richard Hughes
+Yury Umanets
+Mark Hamzy (for some of the early cifs IPv6 work)
+Domen Puncer
+Jesper Juhl (in particular for lots of whitespace/formatting cleanup)
+Vince Negri and Dave Stahl (for finding an important caching bug)
+Adrian Bunk (kcalloc cleanups)
+Miklos Szeredi 
+Kazeon team for various fixes especially for 2.4 version.
+Asser Ferno (Change Notify support)
+Shaggy (Dave Kleikamp) for innumerable small fs suggestions and some good cleanup
+Gunter Kukkukk (testing and suggestions for support of old servers)
+Igor Mammedov (DFS support)
+Jeff Layton (many, many fixes, as well as great work on the cifs Kerberos code)
+Scott Lovenberg
+Pavel Shilovsky (for great work adding SMB2 support, and various SMB3 features)
+Aurelien Aptel (for DFS SMB3 work and some key bug fixes)
+Ronnie Sahlberg (for SMB3 xattr work, bug fixes, and lots of great work on compounding)
+Shirish Pargaonkar (for many ACL patches over the years)
+Sachin Prabhu (many bug fixes, including for reconnect, copy offload and security)
+Paulo Alcantara
+Long Li (some great work on RDMA, SMB Direct)
+
+
+Test case and Bug Report contributors
+-------------------------------------
+Thanks to those in the community who have submitted detailed bug reports
+and debug of problems they have found:  Jochen Dolze, David Blaine,
+Rene Scharfe, Martin Josefsson, Alexander Wild, Anthony Liguori,
+Lars Muller, Urban Widmark, Massimiliano Ferrero, Howard Owen,
+Olaf Kirch, Kieron Briggs, Nick Millington and others. Also special
+mention to the Stanford Checker (SWAT) which pointed out many minor
+bugs in error paths.  Valuable suggestions also have come from Al Viro
+and Dave Miller.
+
+And thanks to the IBM LTC and Power test teams and SuSE and Citrix and RedHat testers for finding multiple bugs during excellent stress test runs.
diff --git a/Documentation/filesystems/cifs/CHANGES b/Documentation/filesystems/cifs/CHANGES
new file mode 100644
index 000000000..1df7f4910
--- /dev/null
+++ b/Documentation/filesystems/cifs/CHANGES
@@ -0,0 +1,4 @@
+See https://wiki.samba.org/index.php/LinuxCIFSKernel for summary
+information (that may be easier to read than parsing the output of
+"git log fs/cifs") about fixes/improvements to CIFS/SMB2/SMB3 support (changes
+to cifs.ko module) by kernel version (and cifs internal module version).
diff --git a/Documentation/filesystems/cifs/README b/Documentation/filesystems/cifs/README
new file mode 100644
index 000000000..4a804619c
--- /dev/null
+++ b/Documentation/filesystems/cifs/README
@@ -0,0 +1,743 @@
+This module supports the SMB3 family of advanced network protocols (as well
+as older dialects, originally called "CIFS" or SMB1).
+
+The CIFS VFS module for Linux supports many advanced network filesystem
+features such as hierarchical DFS like namespace, hardlinks, locking and more.
+It was designed to comply with the SNIA CIFS Technical Reference (which 
+supersedes the 1992 X/Open SMB Standard) as well as to perform best practice 
+practical interoperability with Windows 2000, Windows XP, Samba and equivalent 
+servers.  This code was developed in participation with the Protocol Freedom
+Information Foundation.  CIFS and now SMB3 has now become a defacto
+standard for interoperating between Macs and Windows and major NAS appliances.
+
+Please see
+  MS-SMB2 (for detailed SMB2/SMB3/SMB3.1.1 protocol specification)
+  http://protocolfreedom.org/ and
+  http://samba.org/samba/PFIF/
+for more details.
+
+
+For questions or bug reports please contact:
+    smfrench@gmail.com
+
+See the project page at: https://wiki.samba.org/index.php/LinuxCIFS_utils
+
+Build instructions:
+==================
+For Linux:
+1) Download the kernel (e.g. from http://www.kernel.org)
+and change directory into the top of the kernel directory tree
+(e.g. /usr/src/linux-2.5.73)
+2) make menuconfig (or make xconfig)
+3) select cifs from within the network filesystem choices
+4) save and exit
+5) make
+
+
+Installation instructions:
+=========================
+If you have built the CIFS vfs as module (successfully) simply
+type "make modules_install" (or if you prefer, manually copy the file to
+the modules directory e.g. /lib/modules/2.4.10-4GB/kernel/fs/cifs/cifs.ko).
+
+If you have built the CIFS vfs into the kernel itself, follow the instructions
+for your distribution on how to install a new kernel (usually you
+would simply type "make install").
+
+If you do not have the utility mount.cifs (in the Samba 4.x source tree and on
+the CIFS VFS web site) copy it to the same directory in which mount helpers
+reside (usually /sbin).  Although the helper software is not
+required, mount.cifs is recommended.  Most distros include a "cifs-utils"
+package that includes this utility so it is recommended to install this.
+
+Note that running the Winbind pam/nss module (logon service) on all of your
+Linux clients is useful in mapping Uids and Gids consistently across the
+domain to the proper network user.  The mount.cifs mount helper can be
+found at cifs-utils.git on git.samba.org
+
+If cifs is built as a module, then the size and number of network buffers
+and maximum number of simultaneous requests to one server can be configured.
+Changing these from their defaults is not recommended. By executing modinfo
+	modinfo kernel/fs/cifs/cifs.ko
+on kernel/fs/cifs/cifs.ko the list of configuration changes that can be made
+at module initialization time (by running insmod cifs.ko) can be seen.
+
+Recommendations
+===============
+To improve security the SMB2.1 dialect or later (usually will get SMB3) is now
+the new default. To use old dialects (e.g. to mount Windows XP) use "vers=1.0"
+on mount (or vers=2.0 for Windows Vista).  Note that the CIFS (vers=1.0) is
+much older and less secure than the default dialect SMB3 which includes
+many advanced security features such as downgrade attack detection
+and encrypted shares and stronger signing and authentication algorithms.
+There are additional mount options that may be helpful for SMB3 to get
+improved POSIX behavior (NB: can use vers=3.0 to force only SMB3, never 2.1):
+     "mfsymlinks" and "cifsacl" and "idsfromsid"
+
+Allowing User Mounts
+====================
+To permit users to mount and unmount over directories they own is possible
+with the cifs vfs.  A way to enable such mounting is to mark the mount.cifs
+utility as suid (e.g. "chmod +s /sbin/mount.cifs). To enable users to 
+umount shares they mount requires
+1) mount.cifs version 1.4 or later
+2) an entry for the share in /etc/fstab indicating that a user may
+unmount it e.g.
+//server/usersharename  /mnt/username cifs user 0 0
+
+Note that when the mount.cifs utility is run suid (allowing user mounts), 
+in order to reduce risks, the "nosuid" mount flag is passed in on mount to
+disallow execution of an suid program mounted on the remote target.
+When mount is executed as root, nosuid is not passed in by default,
+and execution of suid programs on the remote target would be enabled
+by default. This can be changed, as with nfs and other filesystems, 
+by simply specifying "nosuid" among the mount options. For user mounts 
+though to be able to pass the suid flag to mount requires rebuilding 
+mount.cifs with the following flag: CIFS_ALLOW_USR_SUID
+
+There is a corresponding manual page for cifs mounting in the Samba 3.0 and
+later source tree in docs/manpages/mount.cifs.8 
+
+Allowing User Unmounts
+======================
+To permit users to ummount directories that they have user mounted (see above),
+the utility umount.cifs may be used.  It may be invoked directly, or if 
+umount.cifs is placed in /sbin, umount can invoke the cifs umount helper
+(at least for most versions of the umount utility) for umount of cifs
+mounts, unless umount is invoked with -i (which will avoid invoking a umount
+helper). As with mount.cifs, to enable user unmounts umount.cifs must be marked
+as suid (e.g. "chmod +s /sbin/umount.cifs") or equivalent (some distributions
+allow adding entries to a file to the /etc/permissions file to achieve the
+equivalent suid effect).  For this utility to succeed the target path
+must be a cifs mount, and the uid of the current user must match the uid
+of the user who mounted the resource.
+
+Also note that the customary way of allowing user mounts and unmounts is 
+(instead of using mount.cifs and unmount.cifs as suid) to add a line
+to the file /etc/fstab for each //server/share you wish to mount, but
+this can become unwieldy when potential mount targets include many
+or  unpredictable UNC names.
+
+Samba Considerations 
+====================
+Most current servers support SMB2.1 and SMB3 which are more secure,
+but there are useful protocol extensions for the older less secure CIFS
+dialect, so to get the maximum benefit if mounting using the older dialect
+(CIFS/SMB1), we recommend using a server that supports the SNIA CIFS
+Unix Extensions standard (e.g. almost any  version of Samba ie version
+2.2.5 or later) but the CIFS vfs works fine with a wide variety of CIFS servers.
+Note that uid, gid and file permissions will display default values if you do 
+not have a server that supports the Unix extensions for CIFS (such as Samba 
+2.2.5 or later).  To enable the Unix CIFS Extensions in the Samba server, add 
+the line: 
+
+	unix extensions = yes
+	
+to your smb.conf file on the server.  Note that the following smb.conf settings 
+are also useful (on the Samba server) when the majority of clients are Unix or 
+Linux: 
+
+	case sensitive = yes
+	delete readonly = yes 
+	ea support = yes
+
+Note that server ea support is required for supporting xattrs from the Linux
+cifs client, and that EA support is present in later versions of Samba (e.g. 
+3.0.6 and later (also EA support works in all versions of Windows, at least to
+shares on NTFS filesystems).  Extended Attribute (xattr) support is an optional
+feature of most Linux filesystems which may require enabling via
+make menuconfig. Client support for extended attributes (user xattr) can be
+disabled on a per-mount basis by specifying "nouser_xattr" on mount.
+
+The CIFS client can get and set POSIX ACLs (getfacl, setfacl) to Samba servers
+version 3.10 and later.  Setting POSIX ACLs requires enabling both XATTR and 
+then POSIX support in the CIFS configuration options when building the cifs
+module.  POSIX ACL support can be disabled on a per mount basic by specifying
+"noacl" on mount.
+ 
+Some administrators may want to change Samba's smb.conf "map archive" and 
+"create mask" parameters from the default.  Unless the create mask is changed
+newly created files can end up with an unnecessarily restrictive default mode,
+which may not be what you want, although if the CIFS Unix extensions are
+enabled on the server and client, subsequent setattr calls (e.g. chmod) can
+fix the mode.  Note that creating special devices (mknod) remotely 
+may require specifying a mkdev function to Samba if you are not using 
+Samba 3.0.6 or later.  For more information on these see the manual pages
+("man smb.conf") on the Samba server system.  Note that the cifs vfs,
+unlike the smbfs vfs, does not read the smb.conf on the client system 
+(the few optional settings are passed in on mount via -o parameters instead).  
+Note that Samba 2.2.7 or later includes a fix that allows the CIFS VFS to delete
+open files (required for strict POSIX compliance).  Windows Servers already 
+supported this feature. Samba server does not allow symlinks that refer to files
+outside of the share, so in Samba versions prior to 3.0.6, most symlinks to
+files with absolute paths (ie beginning with slash) such as:
+	 ln -s /mnt/foo bar
+would be forbidden. Samba 3.0.6 server or later includes the ability to create 
+such symlinks safely by converting unsafe symlinks (ie symlinks to server 
+files that are outside of the share) to a samba specific format on the server
+that is ignored by local server applications and non-cifs clients and that will
+not be traversed by the Samba server).  This is opaque to the Linux client
+application using the cifs vfs. Absolute symlinks will work to Samba 3.0.5 or
+later, but only for remote clients using the CIFS Unix extensions, and will
+be invisbile to Windows clients and typically will not affect local
+applications running on the same server as Samba.  
+
+Use instructions:
+================
+Once the CIFS VFS support is built into the kernel or installed as a module 
+(cifs.ko), you can use mount syntax like the following to access Samba or
+Mac or Windows servers:
+
+  mount -t cifs //9.53.216.11/e$ /mnt -o username=myname,password=mypassword
+
+Before -o the option -v may be specified to make the mount.cifs
+mount helper display the mount steps more verbosely.  
+After -o the following commonly used cifs vfs specific options
+are supported:
+
+  username=<username>
+  password=<password>
+  domain=<domain name>
+  
+Other cifs mount options are described below.  Use of TCP names (in addition to
+ip addresses) is available if the mount helper (mount.cifs) is installed. If
+you do not trust the server to which are mounted, or if you do not have
+cifs signing enabled (and the physical network is insecure), consider use
+of the standard mount options "noexec" and "nosuid" to reduce the risk of 
+running an altered binary on your local system (downloaded from a hostile server
+or altered by a hostile router).
+
+Although mounting using format corresponding to the CIFS URL specification is
+not possible in mount.cifs yet, it is possible to use an alternate format
+for the server and sharename (which is somewhat similar to NFS style mount
+syntax) instead of the more widely used UNC format (i.e. \\server\share):
+  mount -t cifs tcp_name_of_server:share_name /mnt -o user=myname,pass=mypasswd
+
+When using the mount helper mount.cifs, passwords may be specified via alternate
+mechanisms, instead of specifying it after -o using the normal "pass=" syntax
+on the command line:
+1) By including it in a credential file. Specify credentials=filename as one
+of the mount options. Credential files contain two lines
+        username=someuser
+        password=your_password
+2) By specifying the password in the PASSWD environment variable (similarly
+the user name can be taken from the USER environment variable).
+3) By specifying the password in a file by name via PASSWD_FILE
+4) By specifying the password in a file by file descriptor via PASSWD_FD
+
+If no password is provided, mount.cifs will prompt for password entry
+
+Restrictions
+============
+Servers must support either "pure-TCP" (port 445 TCP/IP CIFS connections) or RFC 
+1001/1002 support for "Netbios-Over-TCP/IP." This is not likely to be a 
+problem as most servers support this.
+
+Valid filenames differ between Windows and Linux.  Windows typically restricts
+filenames which contain certain reserved characters (e.g.the character : 
+which is used to delimit the beginning of a stream name by Windows), while
+Linux allows a slightly wider set of valid characters in filenames. Windows
+servers can remap such characters when an explicit mapping is specified in
+the Server's registry.  Samba starting with version 3.10 will allow such 
+filenames (ie those which contain valid Linux characters, which normally
+would be forbidden for Windows/CIFS semantics) as long as the server is
+configured for Unix Extensions (and the client has not disabled
+/proc/fs/cifs/LinuxExtensionsEnabled). In addition the mount option
+"mapposix" can be used on CIFS (vers=1.0) to force the mapping of
+illegal Windows/NTFS/SMB characters to a remap range (this mount parm
+is the default for SMB3). This remap ("mapposix") range is also
+compatible with Mac (and "Services for Mac" on some older Windows).
+
+CIFS VFS Mount Options
+======================
+A partial list of the supported mount options follows:
+  username	The user name to use when trying to establish
+		the CIFS session.
+  password	The user password.  If the mount helper is
+		installed, the user will be prompted for password
+		if not supplied.
+  ip		The ip address of the target server
+  unc		The target server Universal Network Name (export) to 
+		mount.	
+  domain	Set the SMB/CIFS workgroup name prepended to the
+		username during CIFS session establishment
+  forceuid	Set the default uid for inodes to the uid
+		passed in on mount. For mounts to servers
+		which do support the CIFS Unix extensions, such as a
+		properly configured Samba server, the server provides
+		the uid, gid and mode so this parameter should not be
+		specified unless the server and clients uid and gid
+		numbering differ.  If the server and client are in the
+		same domain (e.g. running winbind or nss_ldap) and
+		the server supports the Unix Extensions then the uid
+		and gid can be retrieved from the server (and uid
+		and gid would not have to be specified on the mount.
+		For servers which do not support the CIFS Unix
+		extensions, the default uid (and gid) returned on lookup
+		of existing files will be the uid (gid) of the person
+		who executed the mount (root, except when mount.cifs
+		is configured setuid for user mounts) unless the "uid=" 
+		(gid) mount option is specified. Also note that permission
+		checks (authorization checks) on accesses to a file occur
+		at the server, but there are cases in which an administrator
+		may want to restrict at the client as well.  For those
+		servers which do not report a uid/gid owner
+		(such as Windows), permissions can also be checked at the
+		client, and a crude form of client side permission checking 
+		can be enabled by specifying file_mode and dir_mode on 
+		the client.  (default)
+  forcegid	(similar to above but for the groupid instead of uid) (default)
+  noforceuid	Fill in file owner information (uid) by requesting it from
+		the server if possible. With this option, the value given in
+		the uid= option (on mount) will only be used if the server
+		can not support returning uids on inodes.
+  noforcegid	(similar to above but for the group owner, gid, instead of uid)
+  uid		Set the default uid for inodes, and indicate to the
+		cifs kernel driver which local user mounted. If the server
+		supports the unix extensions the default uid is
+		not used to fill in the owner fields of inodes (files)
+		unless the "forceuid" parameter is specified.
+  gid		Set the default gid for inodes (similar to above).
+  file_mode     If CIFS Unix extensions are not supported by the server
+		this overrides the default mode for file inodes.
+  fsc		Enable local disk caching using FS-Cache (off by default). This
+  		option could be useful to improve performance on a slow link,
+		heavily loaded server and/or network where reading from the
+		disk is faster than reading from the server (over the network).
+		This could also impact scalability positively as the
+		number of calls to the server are reduced. However, local
+		caching is not suitable for all workloads for e.g. read-once
+		type workloads. So, you need to consider carefully your
+		workload/scenario before using this option. Currently, local
+		disk caching is functional for CIFS files opened as read-only.
+  dir_mode      If CIFS Unix extensions are not supported by the server 
+		this overrides the default mode for directory inodes.
+  port		attempt to contact the server on this tcp port, before
+		trying the usual ports (port 445, then 139).
+  iocharset     Codepage used to convert local path names to and from
+		Unicode. Unicode is used by default for network path
+		names if the server supports it.  If iocharset is
+		not specified then the nls_default specified
+		during the local client kernel build will be used.
+		If server does not support Unicode, this parameter is
+		unused.
+  rsize		default read size (usually 16K). The client currently
+		can not use rsize larger than CIFSMaxBufSize. CIFSMaxBufSize
+		defaults to 16K and may be changed (from 8K to the maximum
+		kmalloc size allowed by your kernel) at module install time
+		for cifs.ko. Setting CIFSMaxBufSize to a very large value
+		will cause cifs to use more memory and may reduce performance
+		in some cases.  To use rsize greater than 127K (the original
+		cifs protocol maximum) also requires that the server support
+		a new Unix Capability flag (for very large read) which some
+		newer servers (e.g. Samba 3.0.26 or later) do. rsize can be
+		set from a minimum of 2048 to a maximum of 130048 (127K or
+		CIFSMaxBufSize, whichever is smaller)
+  wsize		default write size (default 57344)
+		maximum wsize currently allowed by CIFS is 57344 (fourteen
+		4096 byte pages)
+  actimeo=n	attribute cache timeout in seconds (default 1 second).
+		After this timeout, the cifs client requests fresh attribute
+		information from the server. This option allows to tune the
+		attribute cache timeout to suit the workload needs. Shorter
+		timeouts mean better the cache coherency, but increased number
+		of calls to the server. Longer timeouts mean reduced number
+		of calls to the server at the expense of less stricter cache
+		coherency checks (i.e. incorrect attribute cache for a short
+		period of time).
+  rw		mount the network share read-write (note that the
+		server may still consider the share read-only)
+  ro		mount network share read-only
+  version	used to distinguish different versions of the
+		mount helper utility (not typically needed)
+  sep		if first mount option (after the -o), overrides
+		the comma as the separator between the mount
+		parms. e.g.
+			-o user=myname,password=mypassword,domain=mydom
+		could be passed instead with period as the separator by
+			-o sep=.user=myname.password=mypassword.domain=mydom
+		this might be useful when comma is contained within username
+		or password or domain. This option is less important
+		when the cifs mount helper cifs.mount (version 1.1 or later)
+		is used.
+  nosuid        Do not allow remote executables with the suid bit 
+		program to be executed.  This is only meaningful for mounts
+		to servers such as Samba which support the CIFS Unix Extensions.
+		If you do not trust the servers in your network (your mount
+		targets) it is recommended that you specify this option for
+		greater security.
+  exec		Permit execution of binaries on the mount.
+  noexec	Do not permit execution of binaries on the mount.
+  dev		Recognize block devices on the remote mount.
+  nodev		Do not recognize devices on the remote mount.
+  suid          Allow remote files on this mountpoint with suid enabled to 
+		be executed (default for mounts when executed as root,
+		nosuid is default for user mounts).
+  credentials   Although ignored by the cifs kernel component, it is used by 
+		the mount helper, mount.cifs. When mount.cifs is installed it
+		opens and reads the credential file specified in order  
+		to obtain the userid and password arguments which are passed to
+		the cifs vfs.
+  guest         Although ignored by the kernel component, the mount.cifs
+		mount helper will not prompt the user for a password
+		if guest is specified on the mount options.  If no
+		password is specified a null password will be used.
+  perm          Client does permission checks (vfs_permission check of uid
+		and gid of the file against the mode and desired operation),
+		Note that this is in addition to the normal ACL check on the
+		target machine done by the server software. 
+		Client permission checking is enabled by default.
+  noperm        Client does not do permission checks.  This can expose
+		files on this mount to access by other users on the local
+		client system. It is typically only needed when the server
+		supports the CIFS Unix Extensions but the UIDs/GIDs on the
+		client and server system do not match closely enough to allow
+		access by the user doing the mount, but it may be useful with
+		non CIFS Unix Extension mounts for cases in which the default
+		mode is specified on the mount but is not to be enforced on the
+		client (e.g. perhaps when MultiUserMount is enabled)
+		Note that this does not affect the normal ACL check on the
+		target machine done by the server software (of the server
+		ACL against the user name provided at mount time).
+  serverino	Use server's inode numbers instead of generating automatically
+		incrementing inode numbers on the client.  Although this will
+		make it easier to spot hardlinked files (as they will have
+		the same inode numbers) and inode numbers may be persistent,
+		note that the server does not guarantee that the inode numbers
+		are unique if multiple server side mounts are exported under a
+		single share (since inode numbers on the servers might not
+		be unique if multiple filesystems are mounted under the same
+		shared higher level directory).  Note that some older
+		(e.g. pre-Windows 2000) do not support returning UniqueIDs
+		or the CIFS Unix Extensions equivalent and for those
+		this mount option will have no effect.  Exporting cifs mounts
+		under nfsd requires this mount option on the cifs mount.
+		This is now the default if server supports the 
+		required network operation.
+  noserverino   Client generates inode numbers (rather than using the actual one
+		from the server). These inode numbers will vary after
+		unmount or reboot which can confuse some applications,
+		but not all server filesystems support unique inode
+		numbers.
+  setuids       If the CIFS Unix extensions are negotiated with the server
+		the client will attempt to set the effective uid and gid of
+		the local process on newly created files, directories, and
+		devices (create, mkdir, mknod).  If the CIFS Unix Extensions
+		are not negotiated, for newly created files and directories
+		instead of using the default uid and gid specified on
+		the mount, cache the new file's uid and gid locally which means
+		that the uid for the file can change when the inode is
+	        reloaded (or the user remounts the share).
+  nosetuids     The client will not attempt to set the uid and gid on
+		on newly created files, directories, and devices (create, 
+		mkdir, mknod) which will result in the server setting the
+		uid and gid to the default (usually the server uid of the
+		user who mounted the share).  Letting the server (rather than
+		the client) set the uid and gid is the default. If the CIFS
+		Unix Extensions are not negotiated then the uid and gid for
+		new files will appear to be the uid (gid) of the mounter or the
+		uid (gid) parameter specified on the mount.
+  netbiosname   When mounting to servers via port 139, specifies the RFC1001
+		source name to use to represent the client netbios machine 
+		name when doing the RFC1001 netbios session initialize.
+  direct        Do not do inode data caching on files opened on this mount.
+		This precludes mmapping files on this mount. In some cases
+		with fast networks and little or no caching benefits on the
+		client (e.g. when the application is doing large sequential
+		reads bigger than page size without rereading the same data) 
+		this can provide better performance than the default
+		behavior which caches reads (readahead) and writes 
+		(writebehind) through the local Linux client pagecache 
+		if oplock (caching token) is granted and held. Note that
+		direct allows write operations larger than page size
+		to be sent to the server.
+  strictcache   Use for switching on strict cache mode. In this mode the
+		client read from the cache all the time it has Oplock Level II,
+		otherwise - read from the server. All written data are stored
+		in the cache, but if the client doesn't have Exclusive Oplock,
+		it writes the data to the server.
+  rwpidforward  Forward pid of a process who opened a file to any read or write
+		operation on that file. This prevent applications like WINE
+		from failing on read and write if we use mandatory brlock style.
+  acl   	Allow setfacl and getfacl to manage posix ACLs if server
+		supports them.  (default)
+  noacl 	Do not allow setfacl and getfacl calls on this mount
+  user_xattr    Allow getting and setting user xattrs (those attributes whose
+		name begins with "user." or "os2.") as OS/2 EAs (extended
+		attributes) to the server.  This allows support of the
+		setfattr and getfattr utilities. (default)
+  nouser_xattr  Do not allow getfattr/setfattr to get/set/list xattrs 
+  mapchars      Translate six of the seven reserved characters (not backslash)
+			*?<>|:
+		to the remap range (above 0xF000), which also
+		allows the CIFS client to recognize files created with
+		such characters by Windows's POSIX emulation. This can
+		also be useful when mounting to most versions of Samba
+		(which also forbids creating and opening files
+		whose names contain any of these seven characters).
+		This has no effect if the server does not support
+		Unicode on the wire.
+ nomapchars     Do not translate any of these seven characters (default).
+ nocase         Request case insensitive path name matching (case
+		sensitive is the default if the server supports it).
+		(mount option "ignorecase" is identical to "nocase")
+ posixpaths     If CIFS Unix extensions are supported, attempt to
+		negotiate posix path name support which allows certain
+		characters forbidden in typical CIFS filenames, without
+		requiring remapping. (default)
+ noposixpaths   If CIFS Unix extensions are supported, do not request
+		posix path name support (this may cause servers to
+		reject creatingfile with certain reserved characters).
+ nounix         Disable the CIFS Unix Extensions for this mount (tree
+		connection). This is rarely needed, but it may be useful
+		in order to turn off multiple settings all at once (ie
+		posix acls, posix locks, posix paths, symlink support
+		and retrieving uids/gids/mode from the server) or to
+		work around a bug in server which implement the Unix
+		Extensions.
+ nobrl          Do not send byte range lock requests to the server.
+		This is necessary for certain applications that break
+		with cifs style mandatory byte range locks (and most
+		cifs servers do not yet support requesting advisory
+		byte range locks).
+ forcemandatorylock Even if the server supports posix (advisory) byte range
+		locking, send only mandatory lock requests.  For some
+		(presumably rare) applications, originally coded for
+		DOS/Windows, which require Windows style mandatory byte range
+		locking, they may be able to take advantage of this option,
+		forcing the cifs client to only send mandatory locks
+		even if the cifs server would support posix advisory locks.
+		"forcemand" is accepted as a shorter form of this mount
+		option.
+ nostrictsync   If this mount option is set, when an application does an
+		fsync call then the cifs client does not send an SMB Flush
+		to the server (to force the server to write all dirty data
+		for this file immediately to disk), although cifs still sends
+		all dirty (cached) file data to the server and waits for the
+		server to respond to the write.  Since SMB Flush can be
+		very slow, and some servers may be reliable enough (to risk
+		delaying slightly flushing the data to disk on the server),
+		turning on this option may be useful to improve performance for
+		applications that fsync too much, at a small risk of server
+		crash.  If this mount option is not set, by default cifs will
+		send an SMB flush request (and wait for a response) on every
+		fsync call.
+ nodfs          Disable DFS (global name space support) even if the
+		server claims to support it.  This can help work around
+		a problem with parsing of DFS paths with Samba server
+		versions 3.0.24 and 3.0.25.
+ remount        remount the share (often used to change from ro to rw mounts
+	        or vice versa)
+ cifsacl        Report mode bits (e.g. on stat) based on the Windows ACL for
+	        the file. (EXPERIMENTAL)
+ servern        Specify the server 's netbios name (RFC1001 name) to use
+		when attempting to setup a session to the server. 
+		This is needed for mounting to some older servers (such
+		as OS/2 or Windows 98 and Windows ME) since they do not
+		support a default server name.  A server name can be up
+		to 15 characters long and is usually uppercased.
+ sfu            When the CIFS Unix Extensions are not negotiated, attempt to
+		create device files and fifos in a format compatible with
+		Services for Unix (SFU).  In addition retrieve bits 10-12
+		of the mode via the SETFILEBITS extended attribute (as
+		SFU does).  In the future the bottom 9 bits of the
+		mode also will be emulated using queries of the security
+		descriptor (ACL).
+ mfsymlinks     Enable support for Minshall+French symlinks
+		(see http://wiki.samba.org/index.php/UNIX_Extensions#Minshall.2BFrench_symlinks)
+		This option is ignored when specified together with the
+		'sfu' option. Minshall+French symlinks are used even if
+		the server supports the CIFS Unix Extensions.
+ sign           Must use packet signing (helps avoid unwanted data modification
+		by intermediate systems in the route).  Note that signing
+		does not work with lanman or plaintext authentication.
+ seal           Must seal (encrypt) all data on this mounted share before
+		sending on the network.  Requires support for Unix Extensions.
+		Note that this differs from the sign mount option in that it
+		causes encryption of data sent over this mounted share but other
+		shares mounted to the same server are unaffected.
+ locallease     This option is rarely needed. Fcntl F_SETLEASE is
+		used by some applications such as Samba and NFSv4 server to
+		check to see whether a file is cacheable.  CIFS has no way
+		to explicitly request a lease, but can check whether a file
+		is cacheable (oplocked).  Unfortunately, even if a file
+		is not oplocked, it could still be cacheable (ie cifs client
+		could grant fcntl leases if no other local processes are using
+		the file) for cases for example such as when the server does not
+		support oplocks and the user is sure that the only updates to
+		the file will be from this client. Specifying this mount option
+		will allow the cifs client to check for leases (only) locally
+		for files which are not oplocked instead of denying leases
+		in that case. (EXPERIMENTAL)
+ sec            Security mode.  Allowed values are:
+			none	attempt to connection as a null user (no name)
+			krb5    Use Kerberos version 5 authentication
+			krb5i   Use Kerberos authentication and packet signing
+			ntlm    Use NTLM password hashing (default)
+			ntlmi   Use NTLM password hashing with signing (if
+				/proc/fs/cifs/PacketSigningEnabled on or if
+				server requires signing also can be the default) 
+			ntlmv2  Use NTLMv2 password hashing      
+			ntlmv2i Use NTLMv2 password hashing with packet signing
+			lanman  (if configured in kernel config) use older
+				lanman hash
+hard		Retry file operations if server is not responding
+soft		Limit retries to unresponsive servers (usually only
+		one retry) before returning an error.  (default)
+
+The mount.cifs mount helper also accepts a few mount options before -o
+including:
+
+	-S      take password from stdin (equivalent to setting the environment
+		variable "PASSWD_FD=0"
+	-V      print mount.cifs version
+	-?      display simple usage information
+
+With most 2.6 kernel versions of modutils, the version of the cifs kernel
+module can be displayed via modinfo.
+
+Misc /proc/fs/cifs Flags and Debug Info
+=======================================
+Informational pseudo-files:
+DebugData		Displays information about active CIFS sessions and
+			shares, features enabled as well as the cifs.ko
+			version.
+Stats			Lists summary resource usage information as well as per
+			share statistics.
+
+Configuration pseudo-files:
+SecurityFlags		Flags which control security negotiation and
+			also packet signing. Authentication (may/must)
+			flags (e.g. for NTLM and/or NTLMv2) may be combined with
+			the signing flags.  Specifying two different password
+			hashing mechanisms (as "must use") on the other hand 
+			does not make much sense. Default flags are 
+				0x07007 
+			(NTLM, NTLMv2 and packet signing allowed).  The maximum 
+			allowable flags if you want to allow mounts to servers
+			using weaker password hashes is 0x37037 (lanman,
+			plaintext, ntlm, ntlmv2, signing allowed).  Some
+			SecurityFlags require the corresponding menuconfig
+			options to be enabled (lanman and plaintext require
+			CONFIG_CIFS_WEAK_PW_HASH for example).  Enabling
+			plaintext authentication currently requires also
+			enabling lanman authentication in the security flags
+			because the cifs module only supports sending
+			laintext passwords using the older lanman dialect
+			form of the session setup SMB.  (e.g. for authentication
+			using plain text passwords, set the SecurityFlags
+			to 0x30030):
+ 
+			may use packet signing 				0x00001
+			must use packet signing				0x01001
+			may use NTLM (most common password hash)	0x00002
+			must use NTLM					0x02002
+			may use NTLMv2					0x00004
+			must use NTLMv2					0x04004
+			may use Kerberos security			0x00008
+			must use Kerberos				0x08008
+			may use lanman (weak) password hash  		0x00010
+			must use lanman password hash			0x10010
+			may use plaintext passwords    			0x00020
+			must use plaintext passwords			0x20020
+			(reserved for future packet encryption)		0x00040
+
+cifsFYI			If set to non-zero value, additional debug information
+			will be logged to the system error log.  This field
+			contains three flags controlling different classes of
+			debugging entries.  The maximum value it can be set
+			to is 7 which enables all debugging points (default 0).
+			Some debugging statements are not compiled into the
+			cifs kernel unless CONFIG_CIFS_DEBUG2 is enabled in the
+			kernel configuration. cifsFYI may be set to one or
+			nore of the following flags (7 sets them all):
+
+			log cifs informational messages			0x01
+			log return codes from cifs entry points		0x02
+			log slow responses (ie which take longer than 1 second)
+			  CONFIG_CIFS_STATS2 must be enabled in .config	0x04
+				
+				
+traceSMB		If set to one, debug information is logged to the
+			system error log with the start of smb requests
+			and responses (default 0)
+LookupCacheEnable	If set to one, inode information is kept cached
+			for one second improving performance of lookups
+			(default 1)
+LinuxExtensionsEnabled	If set to one then the client will attempt to
+			use the CIFS "UNIX" extensions which are optional
+			protocol enhancements that allow CIFS servers
+			to return accurate UID/GID information as well
+			as support symbolic links. If you use servers
+			such as Samba that support the CIFS Unix
+			extensions but do not want to use symbolic link
+			support and want to map the uid and gid fields 
+			to values supplied at mount (rather than the 
+			actual values, then set this to zero. (default 1)
+
+These experimental features and tracing can be enabled by changing flags in 
+/proc/fs/cifs (after the cifs module has been installed or built into the 
+kernel, e.g.  insmod cifs).  To enable a feature set it to 1 e.g.  to enable 
+tracing to the kernel message log type: 
+
+	echo 7 > /proc/fs/cifs/cifsFYI
+	
+cifsFYI functions as a bit mask. Setting it to 1 enables additional kernel
+logging of various informational messages.  2 enables logging of non-zero
+SMB return codes while 4 enables logging of requests that take longer
+than one second to complete (except for byte range lock requests). 
+Setting it to 4 requires CONFIG_CIFS_STATS2 to be set in kernel configuration
+(.config). Setting it to seven enables all three.  Finally, tracing
+the start of smb requests and responses can be enabled via:
+
+	echo 1 > /proc/fs/cifs/traceSMB
+
+Per share (per client mount) statistics are available in /proc/fs/cifs/Stats.
+Additional information is available if CONFIG_CIFS_STATS2 is enabled in the
+kernel configuration (.config).  The statistics returned include counters which
+represent the number of attempted and failed (ie non-zero return code from the
+server) SMB3 (or cifs) requests grouped by request type (read, write, close etc.).
+Also recorded is the total bytes read and bytes written to the server for
+that share.  Note that due to client caching effects this can be less than the
+number of bytes read and written by the application running on the client.
+Statistics can be reset to zero by "echo 0 > /proc/fs/cifs/Stats" which may be
+useful if comparing performance of two different scenarios.
+	
+Also note that "cat /proc/fs/cifs/DebugData" will display information about
+the active sessions and the shares that are mounted.
+
+Enabling Kerberos (extended security) works but requires version 1.2 or later
+of the helper program cifs.upcall to be present and to be configured in the
+/etc/request-key.conf file.  The cifs.upcall helper program is from the Samba
+project(http://www.samba.org). NTLM and NTLMv2 and LANMAN support do not
+require this helper. Note that NTLMv2 security (which does not require the
+cifs.upcall helper program), instead of using Kerberos, is sufficient for
+some use cases.
+
+DFS support allows transparent redirection to shares in an MS-DFS name space.
+In addition, DFS support for target shares which are specified as UNC
+names which begin with host names (rather than IP addresses) requires
+a user space helper (such as cifs.upcall) to be present in order to
+translate host names to ip address, and the user space helper must also
+be configured in the file /etc/request-key.conf.  Samba, Windows servers and
+many NAS appliances support DFS as a way of constructing a global name
+space to ease network configuration and improve reliability.
+
+To use cifs Kerberos and DFS support, the Linux keyutils package should be
+installed and something like the following lines should be added to the
+/etc/request-key.conf file:
+
+create cifs.spnego * * /usr/local/sbin/cifs.upcall %k
+create dns_resolver * * /usr/local/sbin/cifs.upcall %k
+
+CIFS kernel module parameters
+=============================
+These module parameters can be specified or modified either during the time of
+module loading or during the runtime by using the interface
+	/proc/module/cifs/parameters/<param>
+
+i.e. echo "value" > /sys/module/cifs/parameters/<param>
+
+1. enable_oplocks - Enable or disable oplocks. Oplocks are enabled by default.
+		    [Y/y/1]. To disable use any of [N/n/0].
+
diff --git a/Documentation/filesystems/cifs/TODO b/Documentation/filesystems/cifs/TODO
new file mode 100644
index 000000000..852499aed
--- /dev/null
+++ b/Documentation/filesystems/cifs/TODO
@@ -0,0 +1,104 @@
+Version 2.11 September 13, 2017
+
+A Partial List of Missing Features
+==================================
+
+Contributions are welcome.  There are plenty of opportunities
+for visible, important contributions to this module.  Here
+is a partial list of the known problems and missing features:
+
+a) SMB3 (and SMB3.02) missing optional features:
+   - multichannel (started), integration with RDMA
+   - directory leases (improved metadata caching), started (root dir only)
+   - T10 copy offload ie "ODX" (copy chunk, and "Duplicate Extents" ioctl
+     currently the only two server side copy mechanisms supported)
+
+b) improved sparse file support
+
+c) Directory entry caching relies on a 1 second timer, rather than
+using Directory Leases, currently only the root file handle is cached longer
+
+d) quota support (needs minor kernel change since quota calls
+to make it to network filesystems or deviceless filesystems)
+
+e) Compounding (in progress) to reduce number of roundtrips, and also
+better optimize open to reduce redundant opens (using reference counts more).
+
+f) Finish inotify support so kde and gnome file list windows
+will autorefresh (partially complete by Asser). Needs minor kernel
+vfs change to support removing D_NOTIFY on a file.   
+
+g) Add GUI tool to configure /proc/fs/cifs settings and for display of
+the CIFS statistics (started)
+
+h) implement support for security and trusted categories of xattrs
+(requires minor protocol extension) to enable better support for SELINUX
+
+i) Add support for tree connect contexts (see MS-SMB2) a new SMB3.1.1 protocol
+   feature (may be especially useful for virtualization).
+
+j) Create UID mapping facility so server UIDs can be mapped on a per
+mount or a per server basis to client UIDs or nobody if no mapping
+exists. Also better integration with winbind for resolving SID owners
+
+k) Add tools to take advantage of more smb3 specific ioctls and features
+(passthrough ioctl/fsctl for sending various SMB3 fsctls to the server
+is in progress)
+
+l) encrypted file support
+
+m) improved stats gathering, tools (perhaps integration with nfsometer?)
+
+n) allow setting more NTFS/SMB3 file attributes remotely (currently limited to compressed
+file attribute via chflags) and improve user space tools for managing and
+viewing them.
+
+o) mount helper GUI (to simplify the various configuration options on mount)
+
+p) Add support for witness protocol (perhaps ioctl to cifs.ko from user space
+   tool listening on witness protocol RPC) to allow for notification of share
+   move, server failover, and server adapter changes.  And also improve other
+   failover scenarios, e.g. when client knows multiple DFS entries point to
+   different servers, and the server we are connected to has gone down.
+
+q) Allow mount.cifs to be more verbose in reporting errors with dialect
+or unsupported feature errors.
+
+r) updating cifs documentation, and user guide.
+
+s) Addressing bugs found by running a broader set of xfstests in standard
+file system xfstest suite.
+
+t) split cifs and smb3 support into separate modules so legacy (and less
+secure) CIFS dialect can be disabled in environments that don't need it
+and simplify the code.
+
+v) POSIX Extensions for SMB3.1.1 (started, create and mkdir support added
+so far).
+
+KNOWN BUGS
+====================================
+See http://bugzilla.samba.org - search on product "CifsVFS" for
+current bug list.  Also check http://bugzilla.kernel.org (Product = File System, Component = CIFS)
+
+1) existing symbolic links (Windows reparse points) are recognized but
+can not be created remotely. They are implemented for Samba and those that
+support the CIFS Unix extensions, although earlier versions of Samba
+overly restrict the pathnames.
+2) follow_link and readdir code does not follow dfs junctions
+but recognizes them
+
+Misc testing to do
+==================
+1) check out max path names and max path name components against various server
+types. Try nested symlinks (8 deep). Return max path name in stat -f information
+
+2) Improve xfstest's cifs/smb3 enablement and adapt xfstests where needed to test
+cifs/smb3 better
+
+3) Additional performance testing and optimization using iozone and similar - 
+there are some easy changes that can be done to parallelize sequential writes,
+and when signing is disabled to request larger read sizes (larger than 
+negotiated size) and send larger write sizes to modern servers.
+
+4) More exhaustively test against less common servers
diff --git a/Documentation/filesystems/cifs/cifs.txt b/Documentation/filesystems/cifs/cifs.txt
new file mode 100644
index 000000000..677566072
--- /dev/null
+++ b/Documentation/filesystems/cifs/cifs.txt
@@ -0,0 +1,35 @@
+  This is the client VFS module for the SMB3 NAS protocol as well
+  older dialects such as the Common Internet File System (CIFS)
+  protocol which was the successor to the Server Message Block
+  (SMB) protocol, the native file sharing mechanism for most early
+  PC operating systems. New and improved versions of CIFS are now
+  called SMB2 and SMB3. These dialects are also supported by the
+  CIFS VFS module. CIFS is fully supported by network
+  file servers such as Windows 2000, 2003, 2008, 2012 and 2016
+  as well by Samba (which provides excellent CIFS
+  server support for Linux and many other operating systems), Apple
+  systems, as well as most Network Attached Storage vendors, so
+  this network filesystem client can mount to a wide variety of
+  servers.
+
+  The intent of this module is to provide the most advanced network
+  file system function for SMB3 compliant servers, including advanced
+  security features, excellent parallelized high performance i/o, better
+  POSIX compliance, secure per-user session establishment, encryption,
+  high performance safe distributed caching (leases/oplocks), optional packet
+  signing, large files, Unicode support and other internationalization
+  improvements. Since both Samba server and this filesystem client support
+  the CIFS Unix extensions (and in the future SMB3 POSIX extensions),
+  the combination can provide a reasonable alternative to other network and
+  cluster file systems for fileserving in some Linux to Linux environments,
+  not just in Linux to Windows (or Linux to Mac) environments.
+
+  This filesystem has an mount utility (mount.cifs) that can be obtained from
+
+      https://ftp.samba.org/pub/linux-cifs/cifs-utils/
+
+  It must be installed in the directory with the other mount helpers.
+
+  For more information on the module see the project wiki page at
+
+      https://wiki.samba.org/index.php/LinuxCIFS_utils
diff --git a/Documentation/filesystems/cifs/winucase_convert.pl b/Documentation/filesystems/cifs/winucase_convert.pl
new file mode 100755
index 000000000..322a9c833
--- /dev/null
+++ b/Documentation/filesystems/cifs/winucase_convert.pl
@@ -0,0 +1,62 @@
+#!/usr/bin/perl -w
+#
+# winucase_convert.pl -- convert "Windows 8 Upper Case Mapping Table.txt" to
+#                        a two-level set of C arrays.
+#
+#   Copyright 2013: Jeff Layton <jlayton@redhat.com>
+#
+#   This program is free software: you can redistribute it and/or modify
+#   it under the terms of the GNU General Public License as published by
+#   the Free Software Foundation, either version 3 of the License, or
+#   (at your option) any later version.
+#
+#   This program is distributed in the hope that it will be useful,
+#   but WITHOUT ANY WARRANTY; without even the implied warranty of
+#   MERCHANTABILITY or FITNESS FOR A PARTICULAR PURPOSE.  See the
+#   GNU General Public License for more details.
+#
+#   You should have received a copy of the GNU General Public License
+#   along with this program.  If not, see <http://www.gnu.org/licenses/>.
+#
+
+while(<>) {
+	next if (!/^0x(..)(..)\t0x(....)\t/);
+	$firstchar = hex($1);
+	$secondchar = hex($2);
+	$uppercase = hex($3);
+
+	$top[$firstchar][$secondchar] = $uppercase;
+}
+
+for ($i = 0; $i < 256; $i++) {
+	next if (!$top[$i]);
+
+	printf("static const wchar_t t2_%2.2x[256] = {", $i);
+	for ($j = 0; $j < 256; $j++) {
+		if (($j % 8) == 0) {
+			print "\n\t";
+		} else {
+			print " ";
+		}
+		printf("0x%4.4x,", $top[$i][$j] ? $top[$i][$j] : 0);
+	}
+	print "\n};\n\n";
+}
+
+printf("static const wchar_t *const toplevel[256] = {", $i);
+for ($i = 0; $i < 256; $i++) {
+	if (($i % 8) == 0) {
+		print "\n\t";
+	} elsif ($top[$i]) {
+		print " ";
+	} else {
+		print "  ";
+	}
+
+	if ($top[$i]) {
+		printf("t2_%2.2x,", $i);
+	} else {
+		print "NULL,";
+	}
+}
+print "\n};\n\n";
diff --git a/Documentation/filesystems/coda.txt b/Documentation/filesystems/coda.txt
new file mode 100644
index 000000000..613113560
--- /dev/null
+++ b/Documentation/filesystems/coda.txt
@@ -0,0 +1,1673 @@
+NOTE: 
+This is one of the technical documents describing a component of
+Coda -- this document describes the client kernel-Venus interface.
+
+For more information:
+  http://www.coda.cs.cmu.edu
+For user level software needed to run Coda:
+  ftp://ftp.coda.cs.cmu.edu
+
+To run Coda you need to get a user level cache manager for the client,
+named Venus, as well as tools to manipulate ACLs, to log in, etc.  The
+client needs to have the Coda filesystem selected in the kernel
+configuration.
+
+The server needs a user level server and at present does not depend on
+kernel support.
+
+
+
+
+
+
+
+  The Venus kernel interface
+  Peter J. Braam
+  v1.0, Nov 9, 1997
+
+  This document describes the communication between Venus and kernel
+  level filesystem code needed for the operation of the Coda file sys-
+  tem.  This document version is meant to describe the current interface
+  (version 1.0) as well as improvements we envisage.
+  ______________________________________________________________________
+
+  Table of Contents
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+  1. Introduction
+
+  2. Servicing Coda filesystem calls
+
+  3. The message layer
+
+     3.1 Implementation details
+
+  4. The interface at the call level
+
+     4.1 Data structures shared by the kernel and Venus
+     4.2 The pioctl interface
+     4.3 root
+     4.4 lookup
+     4.5 getattr
+     4.6 setattr
+     4.7 access
+     4.8 create
+     4.9 mkdir
+     4.10 link
+     4.11 symlink
+     4.12 remove
+     4.13 rmdir
+     4.14 readlink
+     4.15 open
+     4.16 close
+     4.17 ioctl
+     4.18 rename
+     4.19 readdir
+     4.20 vget
+     4.21 fsync
+     4.22 inactive
+     4.23 rdwr
+     4.24 odymount
+     4.25 ody_lookup
+     4.26 ody_expand
+     4.27 prefetch
+     4.28 signal
+
+  5. The minicache and downcalls
+
+     5.1 INVALIDATE
+     5.2 FLUSH
+     5.3 PURGEUSER
+     5.4 ZAPFILE
+     5.5 ZAPDIR
+     5.6 ZAPVNODE
+     5.7 PURGEFID
+     5.8 REPLACE
+
+  6. Initialization and cleanup
+
+     6.1 Requirements
+
+
+  ______________________________________________________________________
+  0wpage
+
+  11..  IInnttrroodduuccttiioonn
+
+
+
+  A key component in the Coda Distributed File System is the cache
+  manager, _V_e_n_u_s.
+
+
+  When processes on a Coda enabled system access files in the Coda
+  filesystem, requests are directed at the filesystem layer in the
+  operating system. The operating system will communicate with Venus to
+  service the request for the process.  Venus manages a persistent
+  client cache and makes remote procedure calls to Coda file servers and
+  related servers (such as authentication servers) to service these
+  requests it receives from the operating system.  When Venus has
+  serviced a request it replies to the operating system with appropriate
+  return codes, and other data related to the request.  Optionally the
+  kernel support for Coda may maintain a minicache of recently processed
+  requests to limit the number of interactions with Venus.  Venus
+  possesses the facility to inform the kernel when elements from its
+  minicache are no longer valid.
+
+  This document describes precisely this communication between the
+  kernel and Venus.  The definitions of so called upcalls and downcalls
+  will be given with the format of the data they handle. We shall also
+  describe the semantic invariants resulting from the calls.
+
+  Historically Coda was implemented in a BSD file system in Mach 2.6.
+  The interface between the kernel and Venus is very similar to the BSD
+  VFS interface.  Similar functionality is provided, and the format of
+  the parameters and returned data is very similar to the BSD VFS.  This
+  leads to an almost natural environment for implementing a kernel-level
+  filesystem driver for Coda in a BSD system.  However, other operating
+  systems such as Linux and Windows 95 and NT have virtual filesystem
+  with different interfaces.
+
+  To implement Coda on these systems some reverse engineering of the
+  Venus/Kernel protocol is necessary.  Also it came to light that other
+  systems could profit significantly from certain small optimizations
+  and modifications to the protocol. To facilitate this work as well as
+  to make future ports easier, communication between Venus and the
+  kernel should be documented in great detail.  This is the aim of this
+  document.
+
+  0wpage
+
+  22..  SSeerrvviicciinngg CCooddaa ffiilleessyysstteemm ccaallllss
+
+  The service of a request for a Coda file system service originates in
+  a process PP which accessing a Coda file. It makes a system call which
+  traps to the OS kernel. Examples of such calls trapping to the kernel
+  are _r_e_a_d_, _w_r_i_t_e_, _o_p_e_n_, _c_l_o_s_e_, _c_r_e_a_t_e_, _m_k_d_i_r_, _r_m_d_i_r_, _c_h_m_o_d in a Unix
+  context.  Similar calls exist in the Win32 environment, and are named
+  _C_r_e_a_t_e_F_i_l_e_, .
+
+  Generally the operating system handles the request in a virtual
+  filesystem (VFS) layer, which is named I/O Manager in NT and IFS
+  manager in Windows 95.  The VFS is responsible for partial processing
+  of the request and for locating the specific filesystem(s) which will
+  service parts of the request.  Usually the information in the path
+  assists in locating the correct FS drivers.  Sometimes after extensive
+  pre-processing, the VFS starts invoking exported routines in the FS
+  driver.  This is the point where the FS specific processing of the
+  request starts, and here the Coda specific kernel code comes into
+  play.
+
+  The FS layer for Coda must expose and implement several interfaces.
+  First and foremost the VFS must be able to make all necessary calls to
+  the Coda FS layer, so the Coda FS driver must expose the VFS interface
+  as applicable in the operating system. These differ very significantly
+  among operating systems, but share features such as facilities to
+  read/write and create and remove objects.  The Coda FS layer services
+  such VFS requests by invoking one or more well defined services
+  offered by the cache manager Venus.  When the replies from Venus have
+  come back to the FS driver, servicing of the VFS call continues and
+  finishes with a reply to the kernel's VFS. Finally the VFS layer
+  returns to the process.
+
+  As a result of this design a basic interface exposed by the FS driver
+  must allow Venus to manage message traffic.  In particular Venus must
+  be able to retrieve and place messages and to be notified of the
+  arrival of a new message. The notification must be through a mechanism
+  which does not block Venus since Venus must attend to other tasks even
+  when no messages are waiting or being processed.
+
+
+
+
+
+
+                     Interfaces of the Coda FS Driver
+
+  Furthermore the FS layer provides for a special path of communication
+  between a user process and Venus, called the pioctl interface. The
+  pioctl interface is used for Coda specific services, such as
+  requesting detailed information about the persistent cache managed by
+  Venus. Here the involvement of the kernel is minimal.  It identifies
+  the calling process and passes the information on to Venus.  When
+  Venus replies the response is passed back to the caller in unmodified
+  form.
+
+  Finally Venus allows the kernel FS driver to cache the results from
+  certain services.  This is done to avoid excessive context switches
+  and results in an efficient system.  However, Venus may acquire
+  information, for example from the network which implies that cached
+  information must be flushed or replaced. Venus then makes a downcall
+  to the Coda FS layer to request flushes or updates in the cache.  The
+  kernel FS driver handles such requests synchronously.
+
+  Among these interfaces the VFS interface and the facility to place,
+  receive and be notified of messages are platform specific.  We will
+  not go into the calls exported to the VFS layer but we will state the
+  requirements of the message exchange mechanism.
+
+  0wpage
+
+  33..  TThhee mmeessssaaggee llaayyeerr
+
+
+
+  At the lowest level the communication between Venus and the FS driver
+  proceeds through messages.  The synchronization between processes
+  requesting Coda file service and Venus relies on blocking and waking
+  up processes.  The Coda FS driver processes VFS- and pioctl-requests
+  on behalf of a process P, creates messages for Venus, awaits replies
+  and finally returns to the caller.  The implementation of the exchange
+  of messages is platform specific, but the semantics have (so far)
+  appeared to be generally applicable.  Data buffers are created by the
+  FS Driver in kernel memory on behalf of P and copied to user memory in
+  Venus.
+
+  The FS Driver while servicing P makes upcalls to Venus.  Such an
+  upcall is dispatched to Venus by creating a message structure.  The
+  structure contains the identification of P, the message sequence
+  number, the size of the request and a pointer to the data in kernel
+  memory for the request.  Since the data buffer is re-used to hold the
+  reply from Venus, there is a field for the size of the reply.  A flags
+  field is used in the message to precisely record the status of the
+  message.  Additional platform dependent structures involve pointers to
+  determine the position of the message on queues and pointers to
+  synchronization objects.  In the upcall routine the message structure
+  is filled in, flags are set to 0, and it is placed on the _p_e_n_d_i_n_g
+  queue.  The routine calling upcall is responsible for allocating the
+  data buffer; its structure will be described in the next section.
+
+  A facility must exist to notify Venus that the message has been
+  created, and implemented using available synchronization objects in
+  the OS. This notification is done in the upcall context of the process
+  P. When the message is on the pending queue, process P cannot proceed
+  in upcall.  The (kernel mode) processing of P in the filesystem
+  request routine must be suspended until Venus has replied.  Therefore
+  the calling thread in P is blocked in upcall.  A pointer in the
+  message structure will locate the synchronization object on which P is
+  sleeping.
+
+  Venus detects the notification that a message has arrived, and the FS
+  driver allow Venus to retrieve the message with a getmsg_from_kernel
+  call. This action finishes in the kernel by putting the message on the
+  queue of processing messages and setting flags to READ.  Venus is
+  passed the contents of the data buffer. The getmsg_from_kernel call
+  now returns and Venus processes the request.
+
+  At some later point the FS driver receives a message from Venus,
+  namely when Venus calls sendmsg_to_kernel.  At this moment the Coda FS
+  driver looks at the contents of the message and decides if:
+
+
+  +o  the message is a reply for a suspended thread P.  If so it removes
+     the message from the processing queue and marks the message as
+     WRITTEN.  Finally, the FS driver unblocks P (still in the kernel
+     mode context of Venus) and the sendmsg_to_kernel call returns to
+     Venus.  The process P will be scheduled at some point and continues
+     processing its upcall with the data buffer replaced with the reply
+     from Venus.
+
+  +o  The message is a _d_o_w_n_c_a_l_l.  A downcall is a request from Venus to
+     the FS Driver. The FS driver processes the request immediately
+     (usually a cache eviction or replacement) and when it finishes
+     sendmsg_to_kernel returns.
+
+  Now P awakes and continues processing upcall.  There are some
+  subtleties to take account of. First P will determine if it was woken
+  up in upcall by a signal from some other source (for example an
+  attempt to terminate P) or as is normally the case by Venus in its
+  sendmsg_to_kernel call.  In the normal case, the upcall routine will
+  deallocate the message structure and return.  The FS routine can proceed
+  with its processing.
+
+
+
+
+
+
+
+                      Sleeping and IPC arrangements
+
+  In case P is woken up by a signal and not by Venus, it will first look
+  at the flags field.  If the message is not yet READ, the process P can
+  handle its signal without notifying Venus.  If Venus has READ, and
+  the request should not be processed, P can send Venus a signal message
+  to indicate that it should disregard the previous message.  Such
+  signals are put in the queue at the head, and read first by Venus.  If
+  the message is already marked as WRITTEN it is too late to stop the
+  processing.  The VFS routine will now continue.  (-- If a VFS request
+  involves more than one upcall, this can lead to complicated state, an
+  extra field "handle_signals" could be added in the message structure
+  to indicate points of no return have been passed.--)
+
+
+
+  33..11..  IImmpplleemmeennttaattiioonn ddeettaaiillss
+
+  The Unix implementation of this mechanism has been through the
+  implementation of a character device associated with Coda.  Venus
+  retrieves messages by doing a read on the device, replies are sent
+  with a write and notification is through the select system call on the
+  file descriptor for the device.  The process P is kept waiting on an
+  interruptible wait queue object.
+
+  In Windows NT and the DPMI Windows 95 implementation a DeviceIoControl
+  call is used.  The DeviceIoControl call is designed to copy buffers
+  from user memory to kernel memory with OPCODES. The sendmsg_to_kernel
+  is issued as a synchronous call, while the getmsg_from_kernel call is
+  asynchronous.  Windows EventObjects are used for notification of
+  message arrival.  The process P is kept waiting on a KernelEvent
+  object in NT and a semaphore in Windows 95.
+
+  0wpage
+
+  44..  TThhee iinntteerrffaaccee aatt tthhee ccaallll lleevveell
+
+
+  This section describes the upcalls a Coda FS driver can make to Venus.
+  Each of these upcalls make use of two structures: inputArgs and
+  outputArgs.   In pseudo BNF form the structures take the following
+  form:
+
+
+  struct inputArgs {
+      u_long opcode;
+      u_long unique;     /* Keep multiple outstanding msgs distinct */
+      u_short pid;                 /* Common to all */
+      u_short pgid;                /* Common to all */
+      struct CodaCred cred;        /* Common to all */
+
+      <union "in" of call dependent parts of inputArgs>
+  };
+
+  struct outputArgs {
+      u_long opcode;
+      u_long unique;       /* Keep multiple outstanding msgs distinct */
+      u_long result;
+
+      <union "out" of call dependent parts of inputArgs>
+  };
+
+
+
+  Before going on let us elucidate the role of the various fields. The
+  inputArgs start with the opcode which defines the type of service
+  requested from Venus. There are approximately 30 upcalls at present
+  which we will discuss.   The unique field labels the inputArg with a
+  unique number which will identify the message uniquely.  A process and
+  process group id are passed.  Finally the credentials of the caller
+  are included.
+
+  Before delving into the specific calls we need to discuss a variety of
+  data structures shared by the kernel and Venus.
+
+
+
+
+  44..11..  DDaattaa ssttrruuccttuurreess sshhaarreedd bbyy tthhee kkeerrnneell aanndd VVeennuuss
+
+
+  The CodaCred structure defines a variety of user and group ids as
+  they are set for the calling process. The vuid_t and guid_t are 32 bit
+  unsigned integers.  It also defines group membership in an array.  On
+  Unix the CodaCred has proven sufficient to implement good security
+  semantics for Coda but the structure may have to undergo modification
+  for the Windows environment when these mature.
+
+  struct CodaCred {
+      vuid_t cr_uid, cr_euid, cr_suid, cr_fsuid; /* Real, effective, set, fs uid*/
+      vgid_t cr_gid, cr_egid, cr_sgid, cr_fsgid; /* same for groups */
+      vgid_t cr_groups[NGROUPS];        /* Group membership for caller */
+  };
+
+
+
+  NNOOTTEE It is questionable if we need CodaCreds in Venus. Finally Venus
+  doesn't know about groups, although it does create files with the
+  default uid/gid.  Perhaps the list of group membership is superfluous.
+
+
+  The next item is the fundamental identifier used to identify Coda
+  files, the ViceFid.  A fid of a file uniquely defines a file or
+  directory in the Coda filesystem within a _c_e_l_l.   (-- A _c_e_l_l is a
+  group of Coda servers acting under the aegis of a single system
+  control machine or SCM. See the Coda Administration manual for a
+  detailed description of the role of the SCM.--)
+
+
+  typedef struct ViceFid {
+      VolumeId Volume;
+      VnodeId Vnode;
+      Unique_t Unique;
+  } ViceFid;
+
+
+
+  Each of the constituent fields: VolumeId, VnodeId and Unique_t are
+  unsigned 32 bit integers.  We envisage that a further field will need
+  to be prefixed to identify the Coda cell; this will probably take the
+  form of a Ipv6 size IP address naming the Coda cell through DNS.
+
+  The next important structure shared between Venus and the kernel is
+  the attributes of the file.  The following structure is used to
+  exchange information.  It has room for future extensions such as
+  support for device files (currently not present in Coda).
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+
+  struct coda_vattr {
+          enum coda_vtype va_type;        /* vnode type (for create) */
+          u_short         va_mode;        /* files access mode and type */
+          short           va_nlink;       /* number of references to file */
+          vuid_t          va_uid;         /* owner user id */
+          vgid_t          va_gid;         /* owner group id */
+          long            va_fsid;        /* file system id (dev for now) */
+          long            va_fileid;      /* file id */
+          u_quad_t        va_size;        /* file size in bytes */
+          long            va_blocksize;   /* blocksize preferred for i/o */
+          struct timespec va_atime;       /* time of last access */
+          struct timespec va_mtime;       /* time of last modification */
+          struct timespec va_ctime;       /* time file changed */
+          u_long          va_gen;         /* generation number of file */
+          u_long          va_flags;       /* flags defined for file */
+          dev_t           va_rdev;        /* device special file represents */
+          u_quad_t        va_bytes;       /* bytes of disk space held by file */
+          u_quad_t        va_filerev;     /* file modification number */
+          u_int           va_vaflags;     /* operations flags, see below */
+          long            va_spare;       /* remain quad aligned */
+  };
+
+
+
+
+  44..22..  TThhee ppiiooccttll iinntteerrffaaccee
+
+
+  Coda specific requests can be made by application through the pioctl
+  interface. The pioctl is implemented as an ordinary ioctl on a
+  fictitious file /coda/.CONTROL.  The pioctl call opens this file, gets
+  a file handle and makes the ioctl call. Finally it closes the file.
+
+  The kernel involvement in this is limited to providing the facility to
+  open and close and pass the ioctl message _a_n_d to verify that a path in
+  the pioctl data buffers is a file in a Coda filesystem.
+
+  The kernel is handed a data packet of the form:
+
+      struct {
+          const char *path;
+          struct ViceIoctl vidata;
+          int follow;
+      } data;
+
+
+
+  where
+
+
+  struct ViceIoctl {
+          caddr_t in, out;        /* Data to be transferred in, or out */
+          short in_size;          /* Size of input buffer <= 2K */
+          short out_size;         /* Maximum size of output buffer, <= 2K */
+  };
+
+
+
+  The path must be a Coda file, otherwise the ioctl upcall will not be
+  made.
+
+  NNOOTTEE  The data structures and code are a mess.  We need to clean this
+  up.
+
+  We now proceed to document the individual calls:
+
+  0wpage
+
+  44..33..  rroooott
+
+
+  AArrgguummeennttss
+
+     iinn empty
+
+     oouutt
+
+                struct cfs_root_out {
+                    ViceFid VFid;
+                } cfs_root;
+
+
+
+  DDeessccrriippttiioonn This call is made to Venus during the initialization of
+  the Coda filesystem. If the result is zero, the cfs_root structure
+  contains the ViceFid of the root of the Coda filesystem. If a non-zero
+  result is generated, its value is a platform dependent error code
+  indicating the difficulty Venus encountered in locating the root of
+  the Coda filesystem.
+
+  0wpage
+
+  44..44..  llooookkuupp
+
+
+  SSuummmmaarryy Find the ViceFid and type of an object in a directory if it
+  exists.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct  cfs_lookup_in {
+                    ViceFid     VFid;
+                    char        *name;          /* Place holder for data. */
+                } cfs_lookup;
+
+
+
+     oouutt
+
+                struct cfs_lookup_out {
+                    ViceFid VFid;
+                    int vtype;
+                } cfs_lookup;
+
+
+
+  DDeessccrriippttiioonn This call is made to determine the ViceFid and filetype of
+  a directory entry.  The directory entry requested carries name name
+  and Venus will search the directory identified by cfs_lookup_in.VFid.
+  The result may indicate that the name does not exist, or that
+  difficulty was encountered in finding it (e.g. due to disconnection).
+  If the result is zero, the field cfs_lookup_out.VFid contains the
+  targets ViceFid and cfs_lookup_out.vtype the coda_vtype giving the
+  type of object the name designates.
+
+  The name of the object is an 8 bit character string of maximum length
+  CFS_MAXNAMLEN, currently set to 256 (including a 0 terminator.)
+
+  It is extremely important to realize that Venus bitwise ors the field
+  cfs_lookup.vtype with CFS_NOCACHE to indicate that the object should
+  not be put in the kernel name cache.
+
+  NNOOTTEE The type of the vtype is currently wrong.  It should be
+  coda_vtype. Linux does not take note of CFS_NOCACHE.  It should.
+
+  0wpage
+
+  44..55..  ggeettaattttrr
+
+
+  SSuummmmaarryy Get the attributes of a file.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_getattr_in {
+                    ViceFid VFid;
+                    struct coda_vattr attr; /* XXXXX */
+                } cfs_getattr;
+
+
+
+     oouutt
+
+                struct cfs_getattr_out {
+                    struct coda_vattr attr;
+                } cfs_getattr;
+
+
+
+  DDeessccrriippttiioonn This call returns the attributes of the file identified by
+  fid.
+
+  EErrrroorrss Errors can occur if the object with fid does not exist, is
+  unaccessible or if the caller does not have permission to fetch
+  attributes.
+
+  NNoottee Many kernel FS drivers (Linux, NT and Windows 95) need to acquire
+  the attributes as well as the Fid for the instantiation of an internal
+  "inode" or "FileHandle".  A significant improvement in performance on
+  such systems could be made by combining the _l_o_o_k_u_p and _g_e_t_a_t_t_r calls
+  both at the Venus/kernel interaction level and at the RPC level.
+
+  The vattr structure included in the input arguments is superfluous and
+  should be removed.
+
+  0wpage
+
+  44..66..  sseettaattttrr
+
+
+  SSuummmmaarryy Set the attributes of a file.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_setattr_in {
+                    ViceFid VFid;
+                    struct coda_vattr attr;
+                } cfs_setattr;
+
+
+
+
+     oouutt
+        empty
+
+  DDeessccrriippttiioonn The structure attr is filled with attributes to be changed
+  in BSD style.  Attributes not to be changed are set to -1, apart from
+  vtype which is set to VNON. Other are set to the value to be assigned.
+  The only attributes which the FS driver may request to change are the
+  mode, owner, groupid, atime, mtime and ctime.  The return value
+  indicates success or failure.
+
+  EErrrroorrss A variety of errors can occur.  The object may not exist, may
+  be inaccessible, or permission may not be granted by Venus.
+
+  0wpage
+
+  44..77..  aacccceessss
+
+
+  SSuummmmaarryy
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_access_in {
+                    ViceFid     VFid;
+                    int flags;
+                } cfs_access;
+
+
+
+     oouutt
+        empty
+
+  DDeessccrriippttiioonn Verify if access to the object identified by VFid for
+  operations described by flags is permitted.  The result indicates if
+  access will be granted.  It is important to remember that Coda uses
+  ACLs to enforce protection and that ultimately the servers, not the
+  clients enforce the security of the system.  The result of this call
+  will depend on whether a _t_o_k_e_n is held by the user.
+
+  EErrrroorrss The object may not exist, or the ACL describing the protection
+  may not be accessible.
+
+  0wpage
+
+  44..88..  ccrreeaattee
+
+
+  SSuummmmaarryy Invoked to create a file
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_create_in {
+                    ViceFid VFid;
+                    struct coda_vattr attr;
+                    int excl;
+                    int mode;
+                    char        *name;          /* Place holder for data. */
+                } cfs_create;
+
+
+
+
+     oouutt
+
+                struct cfs_create_out {
+                    ViceFid VFid;
+                    struct coda_vattr attr;
+                } cfs_create;
+
+
+
+  DDeessccrriippttiioonn  This upcall is invoked to request creation of a file.
+  The file will be created in the directory identified by VFid, its name
+  will be name, and the mode will be mode.  If excl is set an error will
+  be returned if the file already exists.  If the size field in attr is
+  set to zero the file will be truncated.  The uid and gid of the file
+  are set by converting the CodaCred to a uid using a macro CRTOUID
+  (this macro is platform dependent).  Upon success the VFid and
+  attributes of the file are returned.  The Coda FS Driver will normally
+  instantiate a vnode, inode or file handle at kernel level for the new
+  object.
+
+
+  EErrrroorrss A variety of errors can occur. Permissions may be insufficient.
+  If the object exists and is not a file the error EISDIR is returned
+  under Unix.
+
+  NNOOTTEE The packing of parameters is very inefficient and appears to
+  indicate confusion between the system call creat and the VFS operation
+  create. The VFS operation create is only called to create new objects.
+  This create call differs from the Unix one in that it is not invoked
+  to return a file descriptor. The truncate and exclusive options,
+  together with the mode, could simply be part of the mode as it is
+  under Unix.  There should be no flags argument; this is used in open
+  (2) to return a file descriptor for READ or WRITE mode.
+
+  The attributes of the directory should be returned too, since the size
+  and mtime changed.
+
+  0wpage
+
+  44..99..  mmkkddiirr
+
+
+  SSuummmmaarryy Create a new directory.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_mkdir_in {
+                    ViceFid     VFid;
+                    struct coda_vattr attr;
+                    char        *name;          /* Place holder for data. */
+                } cfs_mkdir;
+
+
+
+     oouutt
+
+                struct cfs_mkdir_out {
+                    ViceFid VFid;
+                    struct coda_vattr attr;
+                } cfs_mkdir;
+
+
+
+
+  DDeessccrriippttiioonn This call is similar to create but creates a directory.
+  Only the mode field in the input parameters is used for creation.
+  Upon successful creation, the attr returned contains the attributes of
+  the new directory.
+
+  EErrrroorrss As for create.
+
+  NNOOTTEE The input parameter should be changed to mode instead of
+  attributes.
+
+  The attributes of the parent should be returned since the size and
+  mtime changes.
+
+  0wpage
+
+  44..1100..  lliinnkk
+
+
+  SSuummmmaarryy Create a link to an existing file.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_link_in {
+                    ViceFid sourceFid;          /* cnode to link *to* */
+                    ViceFid destFid;            /* Directory in which to place link */
+                    char        *tname;         /* Place holder for data. */
+                } cfs_link;
+
+
+
+     oouutt
+        empty
+
+  DDeessccrriippttiioonn This call creates a link to the sourceFid in the directory
+  identified by destFid with name tname.  The source must reside in the
+  target's parent, i.e. the source must be have parent destFid, i.e. Coda
+  does not support cross directory hard links.  Only the return value is
+  relevant.  It indicates success or the type of failure.
+
+  EErrrroorrss The usual errors can occur.0wpage
+
+  44..1111..  ssyymmlliinnkk
+
+
+  SSuummmmaarryy create a symbolic link
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_symlink_in {
+                    ViceFid     VFid;          /* Directory to put symlink in */
+                    char        *srcname;
+                    struct coda_vattr attr;
+                    char        *tname;
+                } cfs_symlink;
+
+
+
+     oouutt
+        none
+
+  DDeessccrriippttiioonn Create a symbolic link. The link is to be placed in the
+  directory identified by VFid and named tname.  It should point to the
+  pathname srcname.  The attributes of the newly created object are to
+  be set to attr.
+
+  EErrrroorrss
+
+  NNOOTTEE The attributes of the target directory should be returned since
+  its size changed.
+
+  0wpage
+
+  44..1122..  rreemmoovvee
+
+
+  SSuummmmaarryy Remove a file
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_remove_in {
+                    ViceFid     VFid;
+                    char        *name;          /* Place holder for data. */
+                } cfs_remove;
+
+
+
+     oouutt
+        none
+
+  DDeessccrriippttiioonn  Remove file named cfs_remove_in.name in directory
+  identified by   VFid.
+
+  EErrrroorrss
+
+  NNOOTTEE The attributes of the directory should be returned since its
+  mtime and size may change.
+
+  0wpage
+
+  44..1133..  rrmmddiirr
+
+
+  SSuummmmaarryy Remove a directory
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_rmdir_in {
+                    ViceFid     VFid;
+                    char        *name;          /* Place holder for data. */
+                } cfs_rmdir;
+
+
+
+     oouutt
+        none
+
+  DDeessccrriippttiioonn Remove the directory with name name from the directory
+  identified by VFid.
+
+  EErrrroorrss
+
+  NNOOTTEE The attributes of the parent directory should be returned since
+  its mtime and size may change.
+
+  0wpage
+
+  44..1144..  rreeaaddlliinnkk
+
+
+  SSuummmmaarryy Read the value of a symbolic link.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_readlink_in {
+                    ViceFid VFid;
+                } cfs_readlink;
+
+
+
+     oouutt
+
+                struct cfs_readlink_out {
+                    int count;
+                    caddr_t     data;           /* Place holder for data. */
+                } cfs_readlink;
+
+
+
+  DDeessccrriippttiioonn This routine reads the contents of symbolic link
+  identified by VFid into the buffer data.  The buffer data must be able
+  to hold any name up to CFS_MAXNAMLEN (PATH or NAM??).
+
+  EErrrroorrss No unusual errors.
+
+  0wpage
+
+  44..1155..  ooppeenn
+
+
+  SSuummmmaarryy Open a file.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_open_in {
+                    ViceFid     VFid;
+                    int flags;
+                } cfs_open;
+
+
+
+     oouutt
+
+                struct cfs_open_out {
+                    dev_t       dev;
+                    ino_t       inode;
+                } cfs_open;
+
+
+
+  DDeessccrriippttiioonn  This request asks Venus to place the file identified by
+  VFid in its cache and to note that the calling process wishes to open
+  it with flags as in open(2).  The return value to the kernel differs
+  for Unix and Windows systems.  For Unix systems the Coda FS Driver is
+  informed of the device and inode number of the container file in the
+  fields dev and inode.  For Windows the path of the container file is
+  returned to the kernel.
+  EErrrroorrss
+
+  NNOOTTEE Currently the cfs_open_out structure is not properly adapted to
+  deal with the Windows case.  It might be best to implement two
+  upcalls, one to open aiming at a container file name, the other at a
+  container file inode.
+
+  0wpage
+
+  44..1166..  cclloossee
+
+
+  SSuummmmaarryy Close a file, update it on the servers.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_close_in {
+                    ViceFid     VFid;
+                    int flags;
+                } cfs_close;
+
+
+
+     oouutt
+        none
+
+  DDeessccrriippttiioonn Close the file identified by VFid.
+
+  EErrrroorrss
+
+  NNOOTTEE The flags argument is bogus and not used.  However, Venus' code
+  has room to deal with an execp input field, probably this field should
+  be used to inform Venus that the file was closed but is still memory
+  mapped for execution.  There are comments about fetching versus not
+  fetching the data in Venus vproc_vfscalls.  This seems silly.  If a
+  file is being closed, the data in the container file is to be the new
+  data.  Here again the execp flag might be in play to create confusion:
+  currently Venus might think a file can be flushed from the cache when
+  it is still memory mapped.  This needs to be understood.
+
+  0wpage
+
+  44..1177..  iiooccttll
+
+
+  SSuummmmaarryy Do an ioctl on a file. This includes the pioctl interface.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_ioctl_in {
+                    ViceFid VFid;
+                    int cmd;
+                    int len;
+                    int rwflag;
+                    char *data;                 /* Place holder for data. */
+                } cfs_ioctl;
+
+
+
+     oouutt
+
+
+                struct cfs_ioctl_out {
+                    int len;
+                    caddr_t     data;           /* Place holder for data. */
+                } cfs_ioctl;
+
+
+
+  DDeessccrriippttiioonn Do an ioctl operation on a file.  The command, len and
+  data arguments are filled as usual.  flags is not used by Venus.
+
+  EErrrroorrss
+
+  NNOOTTEE Another bogus parameter.  flags is not used.  What is the
+  business about PREFETCHING in the Venus code?
+
+
+  0wpage
+
+  44..1188..  rreennaammee
+
+
+  SSuummmmaarryy Rename a fid.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_rename_in {
+                    ViceFid     sourceFid;
+                    char        *srcname;
+                    ViceFid destFid;
+                    char        *destname;
+                } cfs_rename;
+
+
+
+     oouutt
+        none
+
+  DDeessccrriippttiioonn  Rename the object with name srcname in directory
+  sourceFid to destname in destFid.   It is important that the names
+  srcname and destname are 0 terminated strings.  Strings in Unix
+  kernels are not always null terminated.
+
+  EErrrroorrss
+
+  0wpage
+
+  44..1199..  rreeaaddddiirr
+
+
+  SSuummmmaarryy Read directory entries.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_readdir_in {
+                    ViceFid     VFid;
+                    int count;
+                    int offset;
+                } cfs_readdir;
+
+
+
+
+     oouutt
+
+                struct cfs_readdir_out {
+                    int size;
+                    caddr_t     data;           /* Place holder for data. */
+                } cfs_readdir;
+
+
+
+  DDeessccrriippttiioonn Read directory entries from VFid starting at offset and
+  read at most count bytes.  Returns the data in data and returns
+  the size in size.
+
+  EErrrroorrss
+
+  NNOOTTEE This call is not used.  Readdir operations exploit container
+  files.  We will re-evaluate this during the directory revamp which is
+  about to take place.
+
+  0wpage
+
+  44..2200..  vvggeett
+
+
+  SSuummmmaarryy instructs Venus to do an FSDB->Get.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_vget_in {
+                    ViceFid VFid;
+                } cfs_vget;
+
+
+
+     oouutt
+
+                struct cfs_vget_out {
+                    ViceFid VFid;
+                    int vtype;
+                } cfs_vget;
+
+
+
+  DDeessccrriippttiioonn This upcall asks Venus to do a get operation on an fsobj
+  labelled by VFid.
+
+  EErrrroorrss
+
+  NNOOTTEE This operation is not used.  However, it is extremely useful
+  since it can be used to deal with read/write memory mapped files.
+  These can be "pinned" in the Venus cache using vget and released with
+  inactive.
+
+  0wpage
+
+  44..2211..  ffssyynncc
+
+
+  SSuummmmaarryy Tell Venus to update the RVM attributes of a file.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_fsync_in {
+                    ViceFid VFid;
+                } cfs_fsync;
+
+
+
+     oouutt
+        none
+
+  DDeessccrriippttiioonn Ask Venus to update RVM attributes of object VFid. This
+  should be called as part of kernel level fsync type calls.  The
+  result indicates if the syncing was successful.
+
+  EErrrroorrss
+
+  NNOOTTEE Linux does not implement this call. It should.
+
+  0wpage
+
+  44..2222..  iinnaaccttiivvee
+
+
+  SSuummmmaarryy Tell Venus a vnode is no longer in use.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_inactive_in {
+                    ViceFid VFid;
+                } cfs_inactive;
+
+
+
+     oouutt
+        none
+
+  DDeessccrriippttiioonn This operation returns EOPNOTSUPP.
+
+  EErrrroorrss
+
+  NNOOTTEE This should perhaps be removed.
+
+  0wpage
+
+  44..2233..  rrddwwrr
+
+
+  SSuummmmaarryy Read or write from a file
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct cfs_rdwr_in {
+                    ViceFid     VFid;
+                    int rwflag;
+                    int count;
+                    int offset;
+                    int ioflag;
+                    caddr_t     data;           /* Place holder for data. */
+                } cfs_rdwr;
+
+
+
+
+     oouutt
+
+                struct cfs_rdwr_out {
+                    int rwflag;
+                    int count;
+                    caddr_t     data;   /* Place holder for data. */
+                } cfs_rdwr;
+
+
+
+  DDeessccrriippttiioonn This upcall asks Venus to read or write from a file.
+
+  EErrrroorrss
+
+  NNOOTTEE It should be removed since it is against the Coda philosophy that
+  read/write operations never reach Venus.  I have been told the
+  operation does not work.  It is not currently used.
+
+
+  0wpage
+
+  44..2244..  ooddyymmoouunntt
+
+
+  SSuummmmaarryy Allows mounting multiple Coda "filesystems" on one Unix mount
+  point.
+
+  AArrgguummeennttss
+
+     iinn
+
+                struct ody_mount_in {
+                    char        *name;          /* Place holder for data. */
+                } ody_mount;
+
+
+
+     oouutt
+
+                struct ody_mount_out {
+                    ViceFid VFid;
+                } ody_mount;
+
+
+
+  DDeessccrriippttiioonn  Asks Venus to return the rootfid of a Coda system named
+  name.  The fid is returned in VFid.
+
+  EErrrroorrss
+
+  NNOOTTEE This call was used by David for dynamic sets.  It should be
+  removed since it causes a jungle of pointers in the VFS mounting area.
+  It is not used by Coda proper.  Call is not implemented by Venus.
+
+  0wpage
+
+  44..2255..  ooddyy__llooookkuupp
+
+
+  SSuummmmaarryy Looks up something.
+
+  AArrgguummeennttss
+
+     iinn irrelevant
+
+
+     oouutt
+        irrelevant
+
+  DDeessccrriippttiioonn
+
+  EErrrroorrss
+
+  NNOOTTEE Gut it. Call is not implemented by Venus.
+
+  0wpage
+
+  44..2266..  ooddyy__eexxppaanndd
+
+
+  SSuummmmaarryy expands something in a dynamic set.
+
+  AArrgguummeennttss
+
+     iinn irrelevant
+
+     oouutt
+        irrelevant
+
+  DDeessccrriippttiioonn
+
+  EErrrroorrss
+
+  NNOOTTEE Gut it.  Call is not implemented by Venus.
+
+  0wpage
+
+  44..2277..  pprreeffeettcchh
+
+
+  SSuummmmaarryy Prefetch a dynamic set.
+
+  AArrgguummeennttss
+
+     iinn Not documented.
+
+     oouutt
+        Not documented.
+
+  DDeessccrriippttiioonn  Venus worker.cc has support for this call, although it is
+  noted that it doesn't work.  Not surprising, since the kernel does not
+  have support for it. (ODY_PREFETCH is not a defined operation).
+
+  EErrrroorrss
+
+  NNOOTTEE Gut it. It isn't working and isn't used by Coda.
+
+
+  0wpage
+
+  44..2288..  ssiiggnnaall
+
+
+  SSuummmmaarryy Send Venus a signal about an upcall.
+
+  AArrgguummeennttss
+
+     iinn none
+
+     oouutt
+        not applicable.
+
+  DDeessccrriippttiioonn  This is an out-of-band upcall to Venus to inform Venus
+  that the calling process received a signal after Venus read the
+  message from the input queue.  Venus is supposed to clean up the
+  operation.
+
+  EErrrroorrss No reply is given.
+
+  NNOOTTEE We need to better understand what Venus needs to clean up and if
+  it is doing this correctly.  Also we need to handle multiple upcall
+  per system call situations correctly.  It would be important to know
+  what state changes in Venus take place after an upcall for which the
+  kernel is responsible for notifying Venus to clean up (e.g. open
+  definitely is such a state change, but many others are maybe not).
+
+  0wpage
+
+  55..  TThhee mmiinniiccaacchhee aanndd ddoowwnnccaallllss
+
+
+  The Coda FS Driver can cache results of lookup and access upcalls, to
+  limit the frequency of upcalls.  Upcalls carry a price since a process
+  context switch needs to take place.  The counterpart of caching the
+  information is that Venus will notify the FS Driver that cached
+  entries must be flushed or renamed.
+
+  The kernel code generally has to maintain a structure which links the
+  internal file handles (called vnodes in BSD, inodes in Linux and
+  FileHandles in Windows) with the ViceFid's which Venus maintains.  The
+  reason is that frequent translations back and forth are needed in
+  order to make upcalls and use the results of upcalls.  Such linking
+  objects are called ccnnooddeess.
+
+  The current minicache implementations have cache entries which record
+  the following:
+
+  1. the name of the file
+
+  2. the cnode of the directory containing the object
+
+  3. a list of CodaCred's for which the lookup is permitted.
+
+  4. the cnode of the object
+
+  The lookup call in the Coda FS Driver may request the cnode of the
+  desired object from the cache, by passing its name, directory and the
+  CodaCred's of the caller.  The cache will return the cnode or indicate
+  that it cannot be found.  The Coda FS Driver must be careful to
+  invalidate cache entries when it modifies or removes objects.
+
+  When Venus obtains information that indicates that cache entries are
+  no longer valid, it will make a downcall to the kernel.  Downcalls are
+  intercepted by the Coda FS Driver and lead to cache invalidations of
+  the kind described below.  The Coda FS Driver does not return an error
+  unless the downcall data could not be read into kernel memory.
+
+
+  55..11..  IINNVVAALLIIDDAATTEE
+
+
+  No information is available on this call.
+
+
+  55..22..  FFLLUUSSHH
+
+
+
+  AArrgguummeennttss None
+
+  SSuummmmaarryy Flush the name cache entirely.
+
+  DDeessccrriippttiioonn Venus issues this call upon startup and when it dies. This
+  is to prevent stale cache information being held.  Some operating
+  systems allow the kernel name cache to be switched off dynamically.
+  When this is done, this downcall is made.
+
+
+  55..33..  PPUURRGGEEUUSSEERR
+
+
+  AArrgguummeennttss
+
+          struct cfs_purgeuser_out {/* CFS_PURGEUSER is a venus->kernel call */
+              struct CodaCred cred;
+          } cfs_purgeuser;
+
+
+
+  DDeessccrriippttiioonn Remove all entries in the cache carrying the Cred.  This
+  call is issued when tokens for a user expire or are flushed.
+
+
+  55..44..  ZZAAPPFFIILLEE
+
+
+  AArrgguummeennttss
+
+          struct cfs_zapfile_out {  /* CFS_ZAPFILE is a venus->kernel call */
+              ViceFid CodaFid;
+          } cfs_zapfile;
+
+
+
+  DDeessccrriippttiioonn Remove all entries which have the (dir vnode, name) pair.
+  This is issued as a result of an invalidation of cached attributes of
+  a vnode.
+
+  NNOOTTEE Call is not named correctly in NetBSD and Mach.  The minicache
+  zapfile routine takes different arguments. Linux does not implement
+  the invalidation of attributes correctly.
+
+
+
+  55..55..  ZZAAPPDDIIRR
+
+
+  AArrgguummeennttss
+
+          struct cfs_zapdir_out {   /* CFS_ZAPDIR is a venus->kernel call */
+              ViceFid CodaFid;
+          } cfs_zapdir;
+
+
+
+  DDeessccrriippttiioonn Remove all entries in the cache lying in a directory
+  CodaFid, and all children of this directory. This call is issued when
+  Venus receives a callback on the directory.
+
+
+  55..66..  ZZAAPPVVNNOODDEE
+
+
+
+  AArrgguummeennttss
+
+          struct cfs_zapvnode_out { /* CFS_ZAPVNODE is a venus->kernel call */
+              struct CodaCred cred;
+              ViceFid VFid;
+          } cfs_zapvnode;
+
+
+
+  DDeessccrriippttiioonn Remove all entries in the cache carrying the cred and VFid
+  as in the arguments. This downcall is probably never issued.
+
+
+  55..77..  PPUURRGGEEFFIIDD
+
+
+  SSuummmmaarryy
+
+  AArrgguummeennttss
+
+          struct cfs_purgefid_out { /* CFS_PURGEFID is a venus->kernel call */
+              ViceFid CodaFid;
+          } cfs_purgefid;
+
+
+
+  DDeessccrriippttiioonn Flush the attribute for the file. If it is a dir (odd
+  vnode), purge its children from the namecache and remove the file from the
+  namecache.
+
+
+
+  55..88..  RREEPPLLAACCEE
+
+
+  SSuummmmaarryy Replace the Fid's for a collection of names.
+
+  AArrgguummeennttss
+
+          struct cfs_replace_out { /* cfs_replace is a venus->kernel call */
+              ViceFid NewFid;
+              ViceFid OldFid;
+          } cfs_replace;
+
+
+
+  DDeessccrriippttiioonn This routine replaces a ViceFid in the name cache with
+  another.  It is added to allow Venus during reintegration to replace
+  locally allocated temp fids while disconnected with global fids even
+  when the reference counts on those fids are not zero.
+
+  0wpage
+
+  66..  IInniittiiaalliizzaattiioonn aanndd cclleeaannuupp
+
+
+  This section gives brief hints as to desirable features for the Coda
+  FS Driver at startup and upon shutdown or Venus failures.  Before
+  entering the discussion it is useful to repeat that the Coda FS Driver
+  maintains the following data:
+
+
+  1. message queues
+
+  2. cnodes
+
+  3. name cache entries
+
+     The name cache entries are entirely private to the driver, so they
+     can easily be manipulated.   The message queues will generally have
+     clear points of initialization and destruction.  The cnodes are
+     much more delicate.  User processes hold reference counts in Coda
+     filesystems and it can be difficult to clean up the cnodes.
+
+  It can expect requests through:
+
+  1. the message subsystem
+
+  2. the VFS layer
+
+  3. pioctl interface
+
+     Currently the _p_i_o_c_t_l passes through the VFS for Coda so we can
+     treat these similarly.
+
+
+  66..11..  RReeqquuiirreemmeennttss
+
+
+  The following requirements should be accommodated:
+
+  1. The message queues should have open and close routines.  On Unix
+     the opening of the character devices are such routines.
+
+  +o  Before opening, no messages can be placed.
+
+  +o  Opening will remove any old messages still pending.
+
+  +o  Close will notify any sleeping processes that their upcall cannot
+     be completed.
+
+  +o  Close will free all memory allocated by the message queues.
+
+
+  2. At open the namecache shall be initialized to empty state.
+
+  3. Before the message queues are open, all VFS operations will fail.
+     Fortunately this can be achieved by making sure than mounting the
+     Coda filesystem cannot succeed before opening.
+
+  4. After closing of the queues, no VFS operations can succeed.  Here
+     one needs to be careful, since a few operations (lookup,
+     read/write, readdir) can proceed without upcalls.  These must be
+     explicitly blocked.
+
+  5. Upon closing the namecache shall be flushed and disabled.
+
+  6. All memory held by cnodes can be freed without relying on upcalls.
+
+  7. Unmounting the file system can be done without relying on upcalls.
+
+  8. Mounting the Coda filesystem should fail gracefully if Venus cannot
+     get the rootfid or the attributes of the rootfid.  The latter is
+     best implemented by Venus fetching these objects before attempting
+     to mount.
+
+  NNOOTTEE  NetBSD in particular but also Linux have not implemented the
+  above requirements fully.  For smooth operation this needs to be
+  corrected.
+
+
+
diff --git a/Documentation/filesystems/conf.py b/Documentation/filesystems/conf.py
new file mode 100644
index 000000000..ea44172af
--- /dev/null
+++ b/Documentation/filesystems/conf.py
@@ -0,0 +1,10 @@
+# -*- coding: utf-8; mode: python -*-
+
+project = "Linux Filesystems API"
+
+tags.add("subproject")
+
+latex_documents = [
+    ('index', 'filesystems.tex', project,
+     'The kernel development community', 'manual'),
+]
diff --git a/Documentation/filesystems/configfs/configfs.txt b/Documentation/filesystems/configfs/configfs.txt
new file mode 100644
index 000000000..3828e8534
--- /dev/null
+++ b/Documentation/filesystems/configfs/configfs.txt
@@ -0,0 +1,508 @@
+
+configfs - Userspace-driven kernel object configuration.
+
+Joel Becker <joel.becker@oracle.com>
+
+Updated: 31 March 2005
+
+Copyright (c) 2005 Oracle Corporation,
+	Joel Becker <joel.becker@oracle.com>
+
+
+[What is configfs?]
+
+configfs is a ram-based filesystem that provides the converse of
+sysfs's functionality.  Where sysfs is a filesystem-based view of
+kernel objects, configfs is a filesystem-based manager of kernel
+objects, or config_items.
+
+With sysfs, an object is created in kernel (for example, when a device
+is discovered) and it is registered with sysfs.  Its attributes then
+appear in sysfs, allowing userspace to read the attributes via
+readdir(3)/read(2).  It may allow some attributes to be modified via
+write(2).  The important point is that the object is created and
+destroyed in kernel, the kernel controls the lifecycle of the sysfs
+representation, and sysfs is merely a window on all this.
+
+A configfs config_item is created via an explicit userspace operation:
+mkdir(2).  It is destroyed via rmdir(2).  The attributes appear at
+mkdir(2) time, and can be read or modified via read(2) and write(2).
+As with sysfs, readdir(3) queries the list of items and/or attributes.
+symlink(2) can be used to group items together.  Unlike sysfs, the
+lifetime of the representation is completely driven by userspace.  The
+kernel modules backing the items must respond to this.
+
+Both sysfs and configfs can and should exist together on the same
+system.  One is not a replacement for the other.
+
+[Using configfs]
+
+configfs can be compiled as a module or into the kernel.  You can access
+it by doing
+
+	mount -t configfs none /config
+
+The configfs tree will be empty unless client modules are also loaded.
+These are modules that register their item types with configfs as
+subsystems.  Once a client subsystem is loaded, it will appear as a
+subdirectory (or more than one) under /config.  Like sysfs, the
+configfs tree is always there, whether mounted on /config or not.
+
+An item is created via mkdir(2).  The item's attributes will also
+appear at this time.  readdir(3) can determine what the attributes are,
+read(2) can query their default values, and write(2) can store new
+values.  Don't mix more than one attribute in one attribute file.
+
+There are two types of configfs attributes:
+
+* Normal attributes, which similar to sysfs attributes, are small ASCII text
+files, with a maximum size of one page (PAGE_SIZE, 4096 on i386).  Preferably
+only one value per file should be used, and the same caveats from sysfs apply.
+Configfs expects write(2) to store the entire buffer at once.  When writing to
+normal configfs attributes, userspace processes should first read the entire
+file, modify the portions they wish to change, and then write the entire
+buffer back.
+
+* Binary attributes, which are somewhat similar to sysfs binary attributes,
+but with a few slight changes to semantics.  The PAGE_SIZE limitation does not
+apply, but the whole binary item must fit in single kernel vmalloc'ed buffer.
+The write(2) calls from user space are buffered, and the attributes'
+write_bin_attribute method will be invoked on the final close, therefore it is
+imperative for user-space to check the return code of close(2) in order to
+verify that the operation finished successfully.
+To avoid a malicious user OOMing the kernel, there's a per-binary attribute
+maximum buffer value.
+
+When an item needs to be destroyed, remove it with rmdir(2).  An
+item cannot be destroyed if any other item has a link to it (via
+symlink(2)).  Links can be removed via unlink(2).
+
+[Configuring FakeNBD: an Example]
+
+Imagine there's a Network Block Device (NBD) driver that allows you to
+access remote block devices.  Call it FakeNBD.  FakeNBD uses configfs
+for its configuration.  Obviously, there will be a nice program that
+sysadmins use to configure FakeNBD, but somehow that program has to tell
+the driver about it.  Here's where configfs comes in.
+
+When the FakeNBD driver is loaded, it registers itself with configfs.
+readdir(3) sees this just fine:
+
+	# ls /config
+	fakenbd
+
+A fakenbd connection can be created with mkdir(2).  The name is
+arbitrary, but likely the tool will make some use of the name.  Perhaps
+it is a uuid or a disk name:
+
+	# mkdir /config/fakenbd/disk1
+	# ls /config/fakenbd/disk1
+	target device rw
+
+The target attribute contains the IP address of the server FakeNBD will
+connect to.  The device attribute is the device on the server.
+Predictably, the rw attribute determines whether the connection is
+read-only or read-write.
+
+	# echo 10.0.0.1 > /config/fakenbd/disk1/target
+	# echo /dev/sda1 > /config/fakenbd/disk1/device
+	# echo 1 > /config/fakenbd/disk1/rw
+
+That's it.  That's all there is.  Now the device is configured, via the
+shell no less.
+
+[Coding With configfs]
+
+Every object in configfs is a config_item.  A config_item reflects an
+object in the subsystem.  It has attributes that match values on that
+object.  configfs handles the filesystem representation of that object
+and its attributes, allowing the subsystem to ignore all but the
+basic show/store interaction.
+
+Items are created and destroyed inside a config_group.  A group is a
+collection of items that share the same attributes and operations.
+Items are created by mkdir(2) and removed by rmdir(2), but configfs
+handles that.  The group has a set of operations to perform these tasks
+
+A subsystem is the top level of a client module.  During initialization,
+the client module registers the subsystem with configfs, the subsystem
+appears as a directory at the top of the configfs filesystem.  A
+subsystem is also a config_group, and can do everything a config_group
+can.
+
+[struct config_item]
+
+	struct config_item {
+		char                    *ci_name;
+		char                    ci_namebuf[UOBJ_NAME_LEN];
+		struct kref             ci_kref;
+		struct list_head        ci_entry;
+		struct config_item      *ci_parent;
+		struct config_group     *ci_group;
+		struct config_item_type *ci_type;
+		struct dentry           *ci_dentry;
+	};
+
+	void config_item_init(struct config_item *);
+	void config_item_init_type_name(struct config_item *,
+					const char *name,
+					struct config_item_type *type);
+	struct config_item *config_item_get(struct config_item *);
+	void config_item_put(struct config_item *);
+
+Generally, struct config_item is embedded in a container structure, a
+structure that actually represents what the subsystem is doing.  The
+config_item portion of that structure is how the object interacts with
+configfs.
+
+Whether statically defined in a source file or created by a parent
+config_group, a config_item must have one of the _init() functions
+called on it.  This initializes the reference count and sets up the
+appropriate fields.
+
+All users of a config_item should have a reference on it via
+config_item_get(), and drop the reference when they are done via
+config_item_put().
+
+By itself, a config_item cannot do much more than appear in configfs.
+Usually a subsystem wants the item to display and/or store attributes,
+among other things.  For that, it needs a type.
+
+[struct config_item_type]
+
+	struct configfs_item_operations {
+		void (*release)(struct config_item *);
+		int (*allow_link)(struct config_item *src,
+				  struct config_item *target);
+		void (*drop_link)(struct config_item *src,
+				 struct config_item *target);
+	};
+
+	struct config_item_type {
+		struct module                           *ct_owner;
+		struct configfs_item_operations         *ct_item_ops;
+		struct configfs_group_operations        *ct_group_ops;
+		struct configfs_attribute               **ct_attrs;
+		struct configfs_bin_attribute		**ct_bin_attrs;
+	};
+
+The most basic function of a config_item_type is to define what
+operations can be performed on a config_item.  All items that have been
+allocated dynamically will need to provide the ct_item_ops->release()
+method.  This method is called when the config_item's reference count
+reaches zero.
+
+[struct configfs_attribute]
+
+	struct configfs_attribute {
+		char                    *ca_name;
+		struct module           *ca_owner;
+		umode_t                  ca_mode;
+		ssize_t (*show)(struct config_item *, char *);
+		ssize_t (*store)(struct config_item *, const char *, size_t);
+	};
+
+When a config_item wants an attribute to appear as a file in the item's
+configfs directory, it must define a configfs_attribute describing it.
+It then adds the attribute to the NULL-terminated array
+config_item_type->ct_attrs.  When the item appears in configfs, the
+attribute file will appear with the configfs_attribute->ca_name
+filename.  configfs_attribute->ca_mode specifies the file permissions.
+
+If an attribute is readable and provides a ->show method, that method will
+be called whenever userspace asks for a read(2) on the attribute.  If an
+attribute is writable and provides a ->store  method, that method will be
+be called whenever userspace asks for a write(2) on the attribute.
+
+[struct configfs_bin_attribute]
+
+	struct configfs_attribute {
+		struct configfs_attribute	cb_attr;
+		void				*cb_private;
+		size_t				cb_max_size;
+	};
+
+The binary attribute is used when the one needs to use binary blob to
+appear as the contents of a file in the item's configfs directory.
+To do so add the binary attribute to the NULL-terminated array
+config_item_type->ct_bin_attrs, and the item appears in configfs, the
+attribute file will appear with the configfs_bin_attribute->cb_attr.ca_name
+filename.  configfs_bin_attribute->cb_attr.ca_mode specifies the file
+permissions.
+The cb_private member is provided for use by the driver, while the
+cb_max_size member specifies the maximum amount of vmalloc buffer
+to be used.
+
+If binary attribute is readable and the config_item provides a
+ct_item_ops->read_bin_attribute() method, that method will be called
+whenever userspace asks for a read(2) on the attribute.  The converse
+will happen for write(2). The reads/writes are bufferred so only a
+single read/write will occur; the attributes' need not concern itself
+with it.
+
+[struct config_group]
+
+A config_item cannot live in a vacuum.  The only way one can be created
+is via mkdir(2) on a config_group.  This will trigger creation of a
+child item.
+
+	struct config_group {
+		struct config_item		cg_item;
+		struct list_head		cg_children;
+		struct configfs_subsystem 	*cg_subsys;
+		struct list_head		default_groups;
+		struct list_head		group_entry;
+	};
+
+	void config_group_init(struct config_group *group);
+	void config_group_init_type_name(struct config_group *group,
+					 const char *name,
+					 struct config_item_type *type);
+
+
+The config_group structure contains a config_item.  Properly configuring
+that item means that a group can behave as an item in its own right.
+However, it can do more: it can create child items or groups.  This is
+accomplished via the group operations specified on the group's
+config_item_type.
+
+	struct configfs_group_operations {
+		struct config_item *(*make_item)(struct config_group *group,
+						 const char *name);
+		struct config_group *(*make_group)(struct config_group *group,
+						   const char *name);
+		int (*commit_item)(struct config_item *item);
+		void (*disconnect_notify)(struct config_group *group,
+					  struct config_item *item);
+		void (*drop_item)(struct config_group *group,
+				  struct config_item *item);
+	};
+
+A group creates child items by providing the
+ct_group_ops->make_item() method.  If provided, this method is called from mkdir(2) in the group's directory.  The subsystem allocates a new
+config_item (or more likely, its container structure), initializes it,
+and returns it to configfs.  Configfs will then populate the filesystem
+tree to reflect the new item.
+
+If the subsystem wants the child to be a group itself, the subsystem
+provides ct_group_ops->make_group().  Everything else behaves the same,
+using the group _init() functions on the group.
+
+Finally, when userspace calls rmdir(2) on the item or group,
+ct_group_ops->drop_item() is called.  As a config_group is also a
+config_item, it is not necessary for a separate drop_group() method.
+The subsystem must config_item_put() the reference that was initialized
+upon item allocation.  If a subsystem has no work to do, it may omit
+the ct_group_ops->drop_item() method, and configfs will call
+config_item_put() on the item on behalf of the subsystem.
+
+IMPORTANT: drop_item() is void, and as such cannot fail.  When rmdir(2)
+is called, configfs WILL remove the item from the filesystem tree
+(assuming that it has no children to keep it busy).  The subsystem is
+responsible for responding to this.  If the subsystem has references to
+the item in other threads, the memory is safe.  It may take some time
+for the item to actually disappear from the subsystem's usage.  But it
+is gone from configfs.
+
+When drop_item() is called, the item's linkage has already been torn
+down.  It no longer has a reference on its parent and has no place in
+the item hierarchy.  If a client needs to do some cleanup before this
+teardown happens, the subsystem can implement the
+ct_group_ops->disconnect_notify() method.  The method is called after
+configfs has removed the item from the filesystem view but before the
+item is removed from its parent group.  Like drop_item(),
+disconnect_notify() is void and cannot fail.  Client subsystems should
+not drop any references here, as they still must do it in drop_item().
+
+A config_group cannot be removed while it still has child items.  This
+is implemented in the configfs rmdir(2) code.  ->drop_item() will not be
+called, as the item has not been dropped.  rmdir(2) will fail, as the
+directory is not empty.
+
+[struct configfs_subsystem]
+
+A subsystem must register itself, usually at module_init time.  This
+tells configfs to make the subsystem appear in the file tree.
+
+	struct configfs_subsystem {
+		struct config_group	su_group;
+		struct mutex		su_mutex;
+	};
+
+	int configfs_register_subsystem(struct configfs_subsystem *subsys);
+	void configfs_unregister_subsystem(struct configfs_subsystem *subsys);
+
+	A subsystem consists of a toplevel config_group and a mutex.
+The group is where child config_items are created.  For a subsystem,
+this group is usually defined statically.  Before calling
+configfs_register_subsystem(), the subsystem must have initialized the
+group via the usual group _init() functions, and it must also have
+initialized the mutex.
+	When the register call returns, the subsystem is live, and it
+will be visible via configfs.  At that point, mkdir(2) can be called and
+the subsystem must be ready for it.
+
+[An Example]
+
+The best example of these basic concepts is the simple_children
+subsystem/group and the simple_child item in
+samples/configfs/configfs_sample.c. It shows a trivial object displaying
+and storing an attribute, and a simple group creating and destroying
+these children.
+
+[Hierarchy Navigation and the Subsystem Mutex]
+
+There is an extra bonus that configfs provides.  The config_groups and
+config_items are arranged in a hierarchy due to the fact that they
+appear in a filesystem.  A subsystem is NEVER to touch the filesystem
+parts, but the subsystem might be interested in this hierarchy.  For
+this reason, the hierarchy is mirrored via the config_group->cg_children
+and config_item->ci_parent structure members.
+
+A subsystem can navigate the cg_children list and the ci_parent pointer
+to see the tree created by the subsystem.  This can race with configfs'
+management of the hierarchy, so configfs uses the subsystem mutex to
+protect modifications.  Whenever a subsystem wants to navigate the
+hierarchy, it must do so under the protection of the subsystem
+mutex.
+
+A subsystem will be prevented from acquiring the mutex while a newly
+allocated item has not been linked into this hierarchy.   Similarly, it
+will not be able to acquire the mutex while a dropping item has not
+yet been unlinked.  This means that an item's ci_parent pointer will
+never be NULL while the item is in configfs, and that an item will only
+be in its parent's cg_children list for the same duration.  This allows
+a subsystem to trust ci_parent and cg_children while they hold the
+mutex.
+
+[Item Aggregation Via symlink(2)]
+
+configfs provides a simple group via the group->item parent/child
+relationship.  Often, however, a larger environment requires aggregation
+outside of the parent/child connection.  This is implemented via
+symlink(2).
+
+A config_item may provide the ct_item_ops->allow_link() and
+ct_item_ops->drop_link() methods.  If the ->allow_link() method exists,
+symlink(2) may be called with the config_item as the source of the link.
+These links are only allowed between configfs config_items.  Any
+symlink(2) attempt outside the configfs filesystem will be denied.
+
+When symlink(2) is called, the source config_item's ->allow_link()
+method is called with itself and a target item.  If the source item
+allows linking to target item, it returns 0.  A source item may wish to
+reject a link if it only wants links to a certain type of object (say,
+in its own subsystem).
+
+When unlink(2) is called on the symbolic link, the source item is
+notified via the ->drop_link() method.  Like the ->drop_item() method,
+this is a void function and cannot return failure.  The subsystem is
+responsible for responding to the change.
+
+A config_item cannot be removed while it links to any other item, nor
+can it be removed while an item links to it.  Dangling symlinks are not
+allowed in configfs.
+
+[Automatically Created Subgroups]
+
+A new config_group may want to have two types of child config_items.
+While this could be codified by magic names in ->make_item(), it is much
+more explicit to have a method whereby userspace sees this divergence.
+
+Rather than have a group where some items behave differently than
+others, configfs provides a method whereby one or many subgroups are
+automatically created inside the parent at its creation.  Thus,
+mkdir("parent") results in "parent", "parent/subgroup1", up through
+"parent/subgroupN".  Items of type 1 can now be created in
+"parent/subgroup1", and items of type N can be created in
+"parent/subgroupN".
+
+These automatic subgroups, or default groups, do not preclude other
+children of the parent group.  If ct_group_ops->make_group() exists,
+other child groups can be created on the parent group directly.
+
+A configfs subsystem specifies default groups by adding them using the
+configfs_add_default_group() function to the parent config_group
+structure.  Each added group is populated in the configfs tree at the same
+time as the parent group.  Similarly, they are removed at the same time
+as the parent.  No extra notification is provided.  When a ->drop_item()
+method call notifies the subsystem the parent group is going away, it
+also means every default group child associated with that parent group.
+
+As a consequence of this, default groups cannot be removed directly via
+rmdir(2).  They also are not considered when rmdir(2) on the parent
+group is checking for children.
+
+[Dependent Subsystems]
+
+Sometimes other drivers depend on particular configfs items.  For
+example, ocfs2 mounts depend on a heartbeat region item.  If that
+region item is removed with rmdir(2), the ocfs2 mount must BUG or go
+readonly.  Not happy.
+
+configfs provides two additional API calls: configfs_depend_item() and
+configfs_undepend_item().  A client driver can call
+configfs_depend_item() on an existing item to tell configfs that it is
+depended on.  configfs will then return -EBUSY from rmdir(2) for that
+item.  When the item is no longer depended on, the client driver calls
+configfs_undepend_item() on it.
+
+These API cannot be called underneath any configfs callbacks, as
+they will conflict.  They can block and allocate.  A client driver
+probably shouldn't calling them of its own gumption.  Rather it should
+be providing an API that external subsystems call.
+
+How does this work?  Imagine the ocfs2 mount process.  When it mounts,
+it asks for a heartbeat region item.  This is done via a call into the
+heartbeat code.  Inside the heartbeat code, the region item is looked
+up.  Here, the heartbeat code calls configfs_depend_item().  If it
+succeeds, then heartbeat knows the region is safe to give to ocfs2.
+If it fails, it was being torn down anyway, and heartbeat can gracefully
+pass up an error.
+
+[Committable Items]
+
+NOTE: Committable items are currently unimplemented.
+
+Some config_items cannot have a valid initial state.  That is, no
+default values can be specified for the item's attributes such that the
+item can do its work.  Userspace must configure one or more attributes,
+after which the subsystem can start whatever entity this item
+represents.
+
+Consider the FakeNBD device from above.  Without a target address *and*
+a target device, the subsystem has no idea what block device to import.
+The simple example assumes that the subsystem merely waits until all the
+appropriate attributes are configured, and then connects.  This will,
+indeed, work, but now every attribute store must check if the attributes
+are initialized.  Every attribute store must fire off the connection if
+that condition is met.
+
+Far better would be an explicit action notifying the subsystem that the
+config_item is ready to go.  More importantly, an explicit action allows
+the subsystem to provide feedback as to whether the attributes are
+initialized in a way that makes sense.  configfs provides this as
+committable items.
+
+configfs still uses only normal filesystem operations.  An item is
+committed via rename(2).  The item is moved from a directory where it
+can be modified to a directory where it cannot.
+
+Any group that provides the ct_group_ops->commit_item() method has
+committable items.  When this group appears in configfs, mkdir(2) will
+not work directly in the group.  Instead, the group will have two
+subdirectories: "live" and "pending".  The "live" directory does not
+support mkdir(2) or rmdir(2) either.  It only allows rename(2).  The
+"pending" directory does allow mkdir(2) and rmdir(2).  An item is
+created in the "pending" directory.  Its attributes can be modified at
+will.  Userspace commits the item by renaming it into the "live"
+directory.  At this point, the subsystem receives the ->commit_item()
+callback.  If all required attributes are filled to satisfaction, the
+method returns zero and the item is moved to the "live" directory.
+
+As rmdir(2) does not work in the "live" directory, an item must be
+shutdown, or "uncommitted".  Again, this is done via rename(2), this
+time from the "live" directory back to the "pending" one.  The subsystem
+is notified by the ct_group_ops->uncommit_object() method.
+
+
diff --git a/Documentation/filesystems/cramfs.txt b/Documentation/filesystems/cramfs.txt
new file mode 100644
index 000000000..8e19a53d6
--- /dev/null
+++ b/Documentation/filesystems/cramfs.txt
@@ -0,0 +1,118 @@
+
+	Cramfs - cram a filesystem onto a small ROM
+
+cramfs is designed to be simple and small, and to compress things well. 
+
+It uses the zlib routines to compress a file one page at a time, and
+allows random page access.  The meta-data is not compressed, but is
+expressed in a very terse representation to make it use much less
+diskspace than traditional filesystems. 
+
+You can't write to a cramfs filesystem (making it compressible and
+compact also makes it _very_ hard to update on-the-fly), so you have to
+create the disk image with the "mkcramfs" utility.
+
+
+Usage Notes
+-----------
+
+File sizes are limited to less than 16MB.
+
+Maximum filesystem size is a little over 256MB.  (The last file on the
+filesystem is allowed to extend past 256MB.)
+
+Only the low 8 bits of gid are stored.  The current version of
+mkcramfs simply truncates to 8 bits, which is a potential security
+issue.
+
+Hard links are supported, but hard linked files
+will still have a link count of 1 in the cramfs image.
+
+Cramfs directories have no `.' or `..' entries.  Directories (like
+every other file on cramfs) always have a link count of 1.  (There's
+no need to use -noleaf in `find', btw.)
+
+No timestamps are stored in a cramfs, so these default to the epoch
+(1970 GMT).  Recently-accessed files may have updated timestamps, but
+the update lasts only as long as the inode is cached in memory, after
+which the timestamp reverts to 1970, i.e. moves backwards in time.
+
+Currently, cramfs must be written and read with architectures of the
+same endianness, and can be read only by kernels with PAGE_SIZE
+== 4096.  At least the latter of these is a bug, but it hasn't been
+decided what the best fix is.  For the moment if you have larger pages
+you can just change the #define in mkcramfs.c, so long as you don't
+mind the filesystem becoming unreadable to future kernels.
+
+
+Memory Mapped cramfs image
+--------------------------
+
+The CRAMFS_MTD Kconfig option adds support for loading data directly from
+a physical linear memory range (usually non volatile memory like Flash)
+instead of going through the block device layer. This saves some memory
+since no intermediate buffering is necessary to hold the data before
+decompressing.
+
+And when data blocks are kept uncompressed and properly aligned, they will
+automatically be mapped directly into user space whenever possible providing
+eXecute-In-Place (XIP) from ROM of read-only segments. Data segments mapped
+read-write (hence they have to be copied to RAM) may still be compressed in
+the cramfs image in the same file along with non compressed read-only
+segments. Both MMU and no-MMU systems are supported. This is particularly
+handy for tiny embedded systems with very tight memory constraints.
+
+The location of the cramfs image in memory is system dependent. You must
+know the proper physical address where the cramfs image is located and
+configure an MTD device for it. Also, that MTD device must be supported
+by a map driver that implements the "point" method. Examples of such
+MTD drivers are cfi_cmdset_0001 (Intel/Sharp CFI flash) or physmap
+(Flash device in physical memory map). MTD partitions based on such devices
+are fine too. Then that device should be specified with the "mtd:" prefix
+as the mount device argument. For example, to mount the MTD device named
+"fs_partition" on the /mnt directory:
+
+$ mount -t cramfs mtd:fs_partition /mnt
+
+To boot a kernel with this as root filesystem, suffice to specify
+something like "root=mtd:fs_partition" on the kernel command line.
+
+
+Tools
+-----
+
+A version of mkcramfs that can take advantage of the latest capabilities
+described above can be found here:
+
+https://github.com/npitre/cramfs-tools
+
+
+For /usr/share/magic
+--------------------
+
+0	ulelong	0x28cd3d45	Linux cramfs offset 0
+>4	ulelong	x		size %d
+>8	ulelong	x		flags 0x%x
+>12	ulelong	x		future 0x%x
+>16	string	>\0		signature "%.16s"
+>32	ulelong	x		fsid.crc 0x%x
+>36	ulelong	x		fsid.edition %d
+>40	ulelong	x		fsid.blocks %d
+>44	ulelong	x		fsid.files %d
+>48	string	>\0		name "%.16s"
+512	ulelong	0x28cd3d45	Linux cramfs offset 512
+>516	ulelong	x		size %d
+>520	ulelong	x		flags 0x%x
+>524	ulelong	x		future 0x%x
+>528	string	>\0		signature "%.16s"
+>544	ulelong	x		fsid.crc 0x%x
+>548	ulelong	x		fsid.edition %d
+>552	ulelong	x		fsid.blocks %d
+>556	ulelong	x		fsid.files %d
+>560	string	>\0		name "%.16s"
+
+
+Hacker Notes
+------------
+
+See fs/cramfs/README for filesystem layout and implementation notes.
diff --git a/Documentation/filesystems/dax.txt b/Documentation/filesystems/dax.txt
new file mode 100644
index 000000000..70cb68bed
--- /dev/null
+++ b/Documentation/filesystems/dax.txt
@@ -0,0 +1,132 @@
+Direct Access for files
+-----------------------
+
+Motivation
+----------
+
+The page cache is usually used to buffer reads and writes to files.
+It is also used to provide the pages which are mapped into userspace
+by a call to mmap.
+
+For block devices that are memory-like, the page cache pages would be
+unnecessary copies of the original storage.  The DAX code removes the
+extra copy by performing reads and writes directly to the storage device.
+For file mappings, the storage device is mapped directly into userspace.
+
+
+Usage
+-----
+
+If you have a block device which supports DAX, you can make a filesystem
+on it as usual.  The DAX code currently only supports files with a block
+size equal to your kernel's PAGE_SIZE, so you may need to specify a block
+size when creating the filesystem.  When mounting it, use the "-o dax"
+option on the command line or add 'dax' to the options in /etc/fstab.
+
+
+Implementation Tips for Block Driver Writers
+--------------------------------------------
+
+To support DAX in your block driver, implement the 'direct_access'
+block device operation.  It is used to translate the sector number
+(expressed in units of 512-byte sectors) to a page frame number (pfn)
+that identifies the physical page for the memory.  It also returns a
+kernel virtual address that can be used to access the memory.
+
+The direct_access method takes a 'size' parameter that indicates the
+number of bytes being requested.  The function should return the number
+of bytes that can be contiguously accessed at that offset.  It may also
+return a negative errno if an error occurs.
+
+In order to support this method, the storage must be byte-accessible by
+the CPU at all times.  If your device uses paging techniques to expose
+a large amount of memory through a smaller window, then you cannot
+implement direct_access.  Equally, if your device can occasionally
+stall the CPU for an extended period, you should also not attempt to
+implement direct_access.
+
+These block devices may be used for inspiration:
+- brd: RAM backed block device driver
+- dcssblk: s390 dcss block device driver
+- pmem: NVDIMM persistent memory driver
+
+
+Implementation Tips for Filesystem Writers
+------------------------------------------
+
+Filesystem support consists of
+- adding support to mark inodes as being DAX by setting the S_DAX flag in
+  i_flags
+- implementing ->read_iter and ->write_iter operations which use dax_iomap_rw()
+  when inode has S_DAX flag set
+- implementing an mmap file operation for DAX files which sets the
+  VM_MIXEDMAP and VM_HUGEPAGE flags on the VMA, and setting the vm_ops to
+  include handlers for fault, pmd_fault, page_mkwrite, pfn_mkwrite. These
+  handlers should probably call dax_iomap_fault() passing the appropriate
+  fault size and iomap operations.
+- calling iomap_zero_range() passing appropriate iomap operations instead of
+  block_truncate_page() for DAX files
+- ensuring that there is sufficient locking between reads, writes,
+  truncates and page faults
+
+The iomap handlers for allocating blocks must make sure that allocated blocks
+are zeroed out and converted to written extents before being returned to avoid
+exposure of uninitialized data through mmap.
+
+These filesystems may be used for inspiration:
+- ext2: see Documentation/filesystems/ext2.txt
+- ext4: see Documentation/filesystems/ext4.txt
+- xfs:  see Documentation/filesystems/xfs.txt
+
+
+Handling Media Errors
+---------------------
+
+The libnvdimm subsystem stores a record of known media error locations for
+each pmem block device (in gendisk->badblocks). If we fault at such location,
+or one with a latent error not yet discovered, the application can expect
+to receive a SIGBUS. Libnvdimm also allows clearing of these errors by simply
+writing the affected sectors (through the pmem driver, and if the underlying
+NVDIMM supports the clear_poison DSM defined by ACPI).
+
+Since DAX IO normally doesn't go through the driver/bio path, applications or
+sysadmins have an option to restore the lost data from a prior backup/inbuilt
+redundancy in the following ways:
+
+1. Delete the affected file, and restore from a backup (sysadmin route):
+   This will free the file system blocks that were being used by the file,
+   and the next time they're allocated, they will be zeroed first, which
+   happens through the driver, and will clear bad sectors.
+
+2. Truncate or hole-punch the part of the file that has a bad-block (at least
+   an entire aligned sector has to be hole-punched, but not necessarily an
+   entire filesystem block).
+
+These are the two basic paths that allow DAX filesystems to continue operating
+in the presence of media errors. More robust error recovery mechanisms can be
+built on top of this in the future, for example, involving redundancy/mirroring
+provided at the block layer through DM, or additionally, at the filesystem
+level. These would have to rely on the above two tenets, that error clearing
+can happen either by sending an IO through the driver, or zeroing (also through
+the driver).
+
+
+Shortcomings
+------------
+
+Even if the kernel or its modules are stored on a filesystem that supports
+DAX on a block device that supports DAX, they will still be copied into RAM.
+
+The DAX code does not work correctly on architectures which have virtually
+mapped caches such as ARM, MIPS and SPARC.
+
+Calling get_user_pages() on a range of user memory that has been mmaped
+from a DAX file will fail when there are no 'struct page' to describe
+those pages.  This problem has been addressed in some device drivers
+by adding optional struct page support for pages under the control of
+the driver (see CONFIG_NVDIMM_PFN in drivers/nvdimm for an example of
+how to do this). In the non struct page cases O_DIRECT reads/writes to
+those memory ranges from a non-DAX file will fail (note that O_DIRECT
+reads/writes _of a DAX file_ do work, it is the memory that is being
+accessed that is key here).  Other things that will not work in the
+non struct page case include RDMA, sendfile() and splice().
diff --git a/Documentation/filesystems/debugfs.txt b/Documentation/filesystems/debugfs.txt
new file mode 100644
index 000000000..4f45f7114
--- /dev/null
+++ b/Documentation/filesystems/debugfs.txt
@@ -0,0 +1,231 @@
+Copyright 2009 Jonathan Corbet <corbet@lwn.net>
+
+Debugfs exists as a simple way for kernel developers to make information
+available to user space.  Unlike /proc, which is only meant for information
+about a process, or sysfs, which has strict one-value-per-file rules,
+debugfs has no rules at all.  Developers can put any information they want
+there.  The debugfs filesystem is also intended to not serve as a stable
+ABI to user space; in theory, there are no stability constraints placed on
+files exported there.  The real world is not always so simple, though [1];
+even debugfs interfaces are best designed with the idea that they will need
+to be maintained forever.
+
+Debugfs is typically mounted with a command like:
+
+    mount -t debugfs none /sys/kernel/debug
+
+(Or an equivalent /etc/fstab line).
+The debugfs root directory is accessible only to the root user by
+default. To change access to the tree the "uid", "gid" and "mode" mount
+options can be used.
+
+Note that the debugfs API is exported GPL-only to modules.
+
+Code using debugfs should include <linux/debugfs.h>.  Then, the first order
+of business will be to create at least one directory to hold a set of
+debugfs files:
+
+    struct dentry *debugfs_create_dir(const char *name, struct dentry *parent);
+
+This call, if successful, will make a directory called name underneath the
+indicated parent directory.  If parent is NULL, the directory will be
+created in the debugfs root.  On success, the return value is a struct
+dentry pointer which can be used to create files in the directory (and to
+clean it up at the end).  A NULL return value indicates that something went
+wrong.  If ERR_PTR(-ENODEV) is returned, that is an indication that the
+kernel has been built without debugfs support and none of the functions
+described below will work.
+
+The most general way to create a file within a debugfs directory is with:
+
+    struct dentry *debugfs_create_file(const char *name, umode_t mode,
+				       struct dentry *parent, void *data,
+				       const struct file_operations *fops);
+
+Here, name is the name of the file to create, mode describes the access
+permissions the file should have, parent indicates the directory which
+should hold the file, data will be stored in the i_private field of the
+resulting inode structure, and fops is a set of file operations which
+implement the file's behavior.  At a minimum, the read() and/or write()
+operations should be provided; others can be included as needed.  Again,
+the return value will be a dentry pointer to the created file, NULL for
+error, or ERR_PTR(-ENODEV) if debugfs support is missing.
+
+Create a file with an initial size, the following function can be used
+instead:
+
+    struct dentry *debugfs_create_file_size(const char *name, umode_t mode,
+				struct dentry *parent, void *data,
+				const struct file_operations *fops,
+				loff_t file_size);
+
+file_size is the initial file size. The other parameters are the same
+as the function debugfs_create_file.
+
+In a number of cases, the creation of a set of file operations is not
+actually necessary; the debugfs code provides a number of helper functions
+for simple situations.  Files containing a single integer value can be
+created with any of:
+
+    struct dentry *debugfs_create_u8(const char *name, umode_t mode,
+				     struct dentry *parent, u8 *value);
+    struct dentry *debugfs_create_u16(const char *name, umode_t mode,
+				      struct dentry *parent, u16 *value);
+    struct dentry *debugfs_create_u32(const char *name, umode_t mode,
+				      struct dentry *parent, u32 *value);
+    struct dentry *debugfs_create_u64(const char *name, umode_t mode,
+				      struct dentry *parent, u64 *value);
+
+These files support both reading and writing the given value; if a specific
+file should not be written to, simply set the mode bits accordingly.  The
+values in these files are in decimal; if hexadecimal is more appropriate,
+the following functions can be used instead:
+
+    struct dentry *debugfs_create_x8(const char *name, umode_t mode,
+				     struct dentry *parent, u8 *value);
+    struct dentry *debugfs_create_x16(const char *name, umode_t mode,
+				      struct dentry *parent, u16 *value);
+    struct dentry *debugfs_create_x32(const char *name, umode_t mode,
+				      struct dentry *parent, u32 *value);
+    struct dentry *debugfs_create_x64(const char *name, umode_t mode,
+				      struct dentry *parent, u64 *value);
+
+These functions are useful as long as the developer knows the size of the
+value to be exported.  Some types can have different widths on different
+architectures, though, complicating the situation somewhat.  There is a
+function meant to help out in one special case:
+
+    struct dentry *debugfs_create_size_t(const char *name, umode_t mode,
+				         struct dentry *parent, 
+					 size_t *value);
+
+As might be expected, this function will create a debugfs file to represent
+a variable of type size_t.
+
+Boolean values can be placed in debugfs with:
+
+    struct dentry *debugfs_create_bool(const char *name, umode_t mode,
+				       struct dentry *parent, bool *value);
+
+A read on the resulting file will yield either Y (for non-zero values) or
+N, followed by a newline.  If written to, it will accept either upper- or
+lower-case values, or 1 or 0.  Any other input will be silently ignored.
+
+Also, atomic_t values can be placed in debugfs with:
+
+    struct dentry *debugfs_create_atomic_t(const char *name, umode_t mode,
+				struct dentry *parent, atomic_t *value)
+
+A read of this file will get atomic_t values, and a write of this file
+will set atomic_t values.
+
+Another option is exporting a block of arbitrary binary data, with
+this structure and function:
+
+    struct debugfs_blob_wrapper {
+	void *data;
+	unsigned long size;
+    };
+
+    struct dentry *debugfs_create_blob(const char *name, umode_t mode,
+				       struct dentry *parent,
+				       struct debugfs_blob_wrapper *blob);
+
+A read of this file will return the data pointed to by the
+debugfs_blob_wrapper structure.  Some drivers use "blobs" as a simple way
+to return several lines of (static) formatted text output.  This function
+can be used to export binary information, but there does not appear to be
+any code which does so in the mainline.  Note that all files created with
+debugfs_create_blob() are read-only.
+
+If you want to dump a block of registers (something that happens quite
+often during development, even if little such code reaches mainline.
+Debugfs offers two functions: one to make a registers-only file, and
+another to insert a register block in the middle of another sequential
+file.
+
+    struct debugfs_reg32 {
+	char *name;
+	unsigned long offset;
+    };
+
+    struct debugfs_regset32 {
+	struct debugfs_reg32 *regs;
+	int nregs;
+	void __iomem *base;
+    };
+
+    struct dentry *debugfs_create_regset32(const char *name, umode_t mode,
+				     struct dentry *parent,
+				     struct debugfs_regset32 *regset);
+
+    void debugfs_print_regs32(struct seq_file *s, struct debugfs_reg32 *regs,
+			 int nregs, void __iomem *base, char *prefix);
+
+The "base" argument may be 0, but you may want to build the reg32 array
+using __stringify, and a number of register names (macros) are actually
+byte offsets over a base for the register block.
+
+If you want to dump an u32 array in debugfs, you can create file with:
+
+    struct dentry *debugfs_create_u32_array(const char *name, umode_t mode,
+			struct dentry *parent,
+			u32 *array, u32 elements);
+
+The "array" argument provides data, and the "elements" argument is
+the number of elements in the array. Note: Once array is created its
+size can not be changed.
+
+There is a helper function to create device related seq_file:
+
+   struct dentry *debugfs_create_devm_seqfile(struct device *dev,
+				const char *name,
+				struct dentry *parent,
+				int (*read_fn)(struct seq_file *s,
+					void *data));
+
+The "dev" argument is the device related to this debugfs file, and
+the "read_fn" is a function pointer which to be called to print the
+seq_file content.
+
+There are a couple of other directory-oriented helper functions:
+
+    struct dentry *debugfs_rename(struct dentry *old_dir, 
+    				  struct dentry *old_dentry,
+		                  struct dentry *new_dir, 
+				  const char *new_name);
+
+    struct dentry *debugfs_create_symlink(const char *name, 
+                                          struct dentry *parent,
+				      	  const char *target);
+
+A call to debugfs_rename() will give a new name to an existing debugfs
+file, possibly in a different directory.  The new_name must not exist prior
+to the call; the return value is old_dentry with updated information.
+Symbolic links can be created with debugfs_create_symlink().
+
+There is one important thing that all debugfs users must take into account:
+there is no automatic cleanup of any directories created in debugfs.  If a
+module is unloaded without explicitly removing debugfs entries, the result
+will be a lot of stale pointers and no end of highly antisocial behavior.
+So all debugfs users - at least those which can be built as modules - must
+be prepared to remove all files and directories they create there.  A file
+can be removed with:
+
+    void debugfs_remove(struct dentry *dentry);
+
+The dentry value can be NULL, in which case nothing will be removed.
+
+Once upon a time, debugfs users were required to remember the dentry
+pointer for every debugfs file they created so that all files could be
+cleaned up.  We live in more civilized times now, though, and debugfs users
+can call:
+
+    void debugfs_remove_recursive(struct dentry *dentry);
+
+If this function is passed a pointer for the dentry corresponding to the
+top-level directory, the entire hierarchy below that directory will be
+removed.
+
+Notes:
+	[1] http://lwn.net/Articles/309298/
diff --git a/Documentation/filesystems/devpts.txt b/Documentation/filesystems/devpts.txt
new file mode 100644
index 000000000..9f94fe276
--- /dev/null
+++ b/Documentation/filesystems/devpts.txt
@@ -0,0 +1,26 @@
+Each mount of the devpts filesystem is now distinct such that ptys
+and their indicies allocated in one mount are independent from ptys
+and their indicies in all other mounts.
+
+All mounts of the devpts filesystem now create a /dev/pts/ptmx node
+with permissions 0000.
+
+To retain backwards compatibility the a ptmx device node (aka any node
+created with "mknod name c 5 2") when opened will look for an instance
+of devpts under the name "pts" in the same directory as the ptmx device
+node.
+
+As an option instead of placing a /dev/ptmx device node at /dev/ptmx
+it is possible to place a symlink to /dev/pts/ptmx at /dev/ptmx or
+to bind mount /dev/ptx/ptmx to /dev/ptmx.  If you opt for using
+the devpts filesystem in this manner devpts should be mounted with
+the ptmxmode=0666, or chmod 0666 /dev/pts/ptmx should be called.
+
+Total count of pty pairs in all instances is limited by sysctls:
+kernel.pty.max = 4096		- global limit
+kernel.pty.reserve = 1024	- reserved for filesystems mounted from the initial mount namespace
+kernel.pty.nr			- current count of ptys
+
+Per-instance limit could be set by adding mount option "max=<count>".
+This feature was added in kernel 3.4 together with sysctl kernel.pty.reserve.
+In kernels older than 3.4 sysctl kernel.pty.max works as per-instance limit.
diff --git a/Documentation/filesystems/directory-locking b/Documentation/filesystems/directory-locking
new file mode 100644
index 000000000..4e32cb961
--- /dev/null
+++ b/Documentation/filesystems/directory-locking
@@ -0,0 +1,135 @@
+	Locking scheme used for directory operations is based on two
+kinds of locks - per-inode (->i_rwsem) and per-filesystem
+(->s_vfs_rename_mutex).
+
+	When taking the i_rwsem on multiple non-directory objects, we
+always acquire the locks in order by increasing address.  We'll call
+that "inode pointer" order in the following.
+
+	For our purposes all operations fall in 5 classes:
+
+1) read access.  Locking rules: caller locks directory we are accessing.
+The lock is taken shared.
+
+2) object creation.  Locking rules: same as above, but the lock is taken
+exclusive.
+
+3) object removal.  Locking rules: caller locks parent, finds victim,
+locks victim and calls the method.  Locks are exclusive.
+
+4) rename() that is _not_ cross-directory.  Locking rules: caller locks
+the parent and finds source and target.  In case of exchange (with
+RENAME_EXCHANGE in flags argument) lock both.  In any case,
+if the target already exists, lock it.  If the source is a non-directory,
+lock it.  If we need to lock both, lock them in inode pointer order.
+Then call the method.  All locks are exclusive.
+NB: we might get away with locking the the source (and target in exchange
+case) shared.
+
+5) link creation.  Locking rules:
+	* lock parent
+	* check that source is not a directory
+	* lock source
+	* call the method.
+All locks are exclusive.
+
+6) cross-directory rename.  The trickiest in the whole bunch.  Locking
+rules:
+	* lock the filesystem
+	* lock parents in "ancestors first" order.
+	* find source and target.
+	* if old parent is equal to or is a descendent of target
+		fail with -ENOTEMPTY
+	* if new parent is equal to or is a descendent of source
+		fail with -ELOOP
+	* If it's an exchange, lock both the source and the target.
+	* If the target exists, lock it.  If the source is a non-directory,
+	  lock it.  If we need to lock both, do so in inode pointer order.
+	* call the method.
+All ->i_rwsem are taken exclusive.  Again, we might get away with locking
+the the source (and target in exchange case) shared.
+
+The rules above obviously guarantee that all directories that are going to be
+read, modified or removed by method will be locked by caller.
+
+
+If no directory is its own ancestor, the scheme above is deadlock-free.
+Proof:
+
+	First of all, at any moment we have a partial ordering of the
+objects - A < B iff A is an ancestor of B.
+
+	That ordering can change.  However, the following is true:
+
+(1) if object removal or non-cross-directory rename holds lock on A and
+    attempts to acquire lock on B, A will remain the parent of B until we
+    acquire the lock on B.  (Proof: only cross-directory rename can change
+    the parent of object and it would have to lock the parent).
+
+(2) if cross-directory rename holds the lock on filesystem, order will not
+    change until rename acquires all locks.  (Proof: other cross-directory
+    renames will be blocked on filesystem lock and we don't start changing
+    the order until we had acquired all locks).
+
+(3) locks on non-directory objects are acquired only after locks on
+    directory objects, and are acquired in inode pointer order.
+    (Proof: all operations but renames take lock on at most one
+    non-directory object, except renames, which take locks on source and
+    target in inode pointer order in the case they are not directories.)
+
+	Now consider the minimal deadlock.  Each process is blocked on
+attempt to acquire some lock and already holds at least one lock.  Let's
+consider the set of contended locks.  First of all, filesystem lock is
+not contended, since any process blocked on it is not holding any locks.
+Thus all processes are blocked on ->i_rwsem.
+
+	By (3), any process holding a non-directory lock can only be
+waiting on another non-directory lock with a larger address.  Therefore
+the process holding the "largest" such lock can always make progress, and
+non-directory objects are not included in the set of contended locks.
+
+	Thus link creation can't be a part of deadlock - it can't be
+blocked on source and it means that it doesn't hold any locks.
+
+	Any contended object is either held by cross-directory rename or
+has a child that is also contended.  Indeed, suppose that it is held by
+operation other than cross-directory rename.  Then the lock this operation
+is blocked on belongs to child of that object due to (1).
+
+	It means that one of the operations is cross-directory rename.
+Otherwise the set of contended objects would be infinite - each of them
+would have a contended child and we had assumed that no object is its
+own descendent.  Moreover, there is exactly one cross-directory rename
+(see above).
+
+	Consider the object blocking the cross-directory rename.  One
+of its descendents is locked by cross-directory rename (otherwise we
+would again have an infinite set of contended objects).  But that
+means that cross-directory rename is taking locks out of order.  Due
+to (2) the order hadn't changed since we had acquired filesystem lock.
+But locking rules for cross-directory rename guarantee that we do not
+try to acquire lock on descendent before the lock on ancestor.
+Contradiction.  I.e.  deadlock is impossible.  Q.E.D.
+
+
+	These operations are guaranteed to avoid loop creation.  Indeed,
+the only operation that could introduce loops is cross-directory rename.
+Since the only new (parent, child) pair added by rename() is (new parent,
+source), such loop would have to contain these objects and the rest of it
+would have to exist before rename().  I.e. at the moment of loop creation
+rename() responsible for that would be holding filesystem lock and new parent
+would have to be equal to or a descendent of source.  But that means that
+new parent had been equal to or a descendent of source since the moment when
+we had acquired filesystem lock and rename() would fail with -ELOOP in that
+case.
+
+	While this locking scheme works for arbitrary DAGs, it relies on
+ability to check that directory is a descendent of another object.  Current
+implementation assumes that directory graph is a tree.  This assumption is
+also preserved by all operations (cross-directory rename on a tree that would
+not introduce a cycle will leave it a tree and link() fails for directories).
+
+	Notice that "directory" in the above == "anything that might have
+children", so if we are going to introduce hybrid objects we will need
+either to make sure that link(2) doesn't work for them or to make changes
+in is_subdir() that would make it work even in presence of such beasts.
diff --git a/Documentation/filesystems/dlmfs.txt b/Documentation/filesystems/dlmfs.txt
new file mode 100644
index 000000000..fcf4d509d
--- /dev/null
+++ b/Documentation/filesystems/dlmfs.txt
@@ -0,0 +1,130 @@
+dlmfs
+==================
+A minimal DLM userspace interface implemented via a virtual file
+system.
+
+dlmfs is built with OCFS2 as it requires most of its infrastructure.
+
+Project web page:    http://ocfs2.wiki.kernel.org
+Tools web page:      https://github.com/markfasheh/ocfs2-tools
+OCFS2 mailing lists: http://oss.oracle.com/projects/ocfs2/mailman/
+
+All code copyright 2005 Oracle except when otherwise noted.
+
+CREDITS
+=======
+
+Some code taken from ramfs which is Copyright (C) 2000 Linus Torvalds
+and Transmeta Corp.
+
+Mark Fasheh <mark.fasheh@oracle.com>
+
+Caveats
+=======
+- Right now it only works with the OCFS2 DLM, though support for other
+  DLM implementations should not be a major issue.
+
+Mount options
+=============
+None
+
+Usage
+=====
+
+If you're just interested in OCFS2, then please see ocfs2.txt. The
+rest of this document will be geared towards those who want to use
+dlmfs for easy to setup and easy to use clustered locking in
+userspace.
+
+Setup
+=====
+
+dlmfs requires that the OCFS2 cluster infrastructure be in
+place. Please download ocfs2-tools from the above url and configure a
+cluster.
+
+You'll want to start heartbeating on a volume which all the nodes in
+your lockspace can access. The easiest way to do this is via
+ocfs2_hb_ctl (distributed with ocfs2-tools). Right now it requires
+that an OCFS2 file system be in place so that it can automatically
+find its heartbeat area, though it will eventually support heartbeat
+against raw disks.
+
+Please see the ocfs2_hb_ctl and mkfs.ocfs2 manual pages distributed
+with ocfs2-tools.
+
+Once you're heartbeating, DLM lock 'domains' can be easily created /
+destroyed and locks within them accessed.
+
+Locking
+=======
+
+Users may access dlmfs via standard file system calls, or they can use
+'libo2dlm' (distributed with ocfs2-tools) which abstracts the file
+system calls and presents a more traditional locking api.
+
+dlmfs handles lock caching automatically for the user, so a lock
+request for an already acquired lock will not generate another DLM
+call. Userspace programs are assumed to handle their own local
+locking.
+
+Two levels of locks are supported - Shared Read, and Exclusive.
+Also supported is a Trylock operation.
+
+For information on the libo2dlm interface, please see o2dlm.h,
+distributed with ocfs2-tools.
+
+Lock value blocks can be read and written to a resource via read(2)
+and write(2) against the fd obtained via your open(2) call. The
+maximum currently supported LVB length is 64 bytes (though that is an
+OCFS2 DLM limitation). Through this mechanism, users of dlmfs can share
+small amounts of data amongst their nodes.
+
+mkdir(2) signals dlmfs to join a domain (which will have the same name
+as the resulting directory)
+
+rmdir(2) signals dlmfs to leave the domain
+
+Locks for a given domain are represented by regular inodes inside the
+domain directory.  Locking against them is done via the open(2) system
+call.
+
+The open(2) call will not return until your lock has been granted or
+an error has occurred, unless it has been instructed to do a trylock
+operation. If the lock succeeds, you'll get an fd.
+
+open(2) with O_CREAT to ensure the resource inode is created - dlmfs does
+not automatically create inodes for existing lock resources.
+
+Open Flag     Lock Request Type
+---------     -----------------
+O_RDONLY      Shared Read
+O_RDWR        Exclusive
+
+Open Flag     Resulting Locking Behavior
+---------     --------------------------
+O_NONBLOCK    Trylock operation
+
+You must provide exactly one of O_RDONLY or O_RDWR.
+
+If O_NONBLOCK is also provided and the trylock operation was valid but
+could not lock the resource then open(2) will return ETXTBUSY.
+
+close(2) drops the lock associated with your fd.
+
+Modes passed to mkdir(2) or open(2) are adhered to locally. Chown is
+supported locally as well. This means you can use them to restrict
+access to the resources via dlmfs on your local node only.
+
+The resource LVB may be read from the fd in either Shared Read or
+Exclusive modes via the read(2) system call. It can be written via
+write(2) only when open in Exclusive mode.
+
+Once written, an LVB will be visible to other nodes who obtain Read
+Only or higher level locks on the resource.
+
+See Also
+========
+http://opendlm.sourceforge.net/cvsmirror/opendlm/docs/dlmbook_final.pdf
+
+For more information on the VMS distributed locking API.
diff --git a/Documentation/filesystems/dnotify.txt b/Documentation/filesystems/dnotify.txt
new file mode 100644
index 000000000..15156883d
--- /dev/null
+++ b/Documentation/filesystems/dnotify.txt
@@ -0,0 +1,70 @@
+		Linux Directory Notification
+		============================
+
+	   Stephen Rothwell <sfr@canb.auug.org.au>
+
+The intention of directory notification is to allow user applications
+to be notified when a directory, or any of the files in it, are changed.
+The basic mechanism involves the application registering for notification
+on a directory using a fcntl(2) call and the notifications themselves
+being delivered using signals.
+
+The application decides which "events" it wants to be notified about.
+The currently defined events are:
+
+	DN_ACCESS	A file in the directory was accessed (read)
+	DN_MODIFY	A file in the directory was modified (write,truncate)
+	DN_CREATE	A file was created in the directory
+	DN_DELETE	A file was unlinked from directory
+	DN_RENAME	A file in the directory was renamed
+	DN_ATTRIB	A file in the directory had its attributes
+			changed (chmod,chown)
+
+Usually, the application must reregister after each notification, but
+if DN_MULTISHOT is or'ed with the event mask, then the registration will
+remain until explicitly removed (by registering for no events).
+
+By default, SIGIO will be delivered to the process and no other useful
+information.  However, if the F_SETSIG fcntl(2) call is used to let the
+kernel know which signal to deliver, a siginfo structure will be passed to
+the signal handler and the si_fd member of that structure will contain the
+file descriptor associated with the directory in which the event occurred.
+
+Preferably the application will choose one of the real time signals
+(SIGRTMIN + <n>) so that the notifications may be queued.  This is
+especially important if DN_MULTISHOT is specified.  Note that SIGRTMIN
+is often blocked, so it is better to use (at least) SIGRTMIN + 1.
+
+Implementation expectations (features and bugs :-))
+---------------------------
+
+The notification should work for any local access to files even if the
+actual file system is on a remote server.  This implies that remote
+access to files served by local user mode servers should be notified.
+Also, remote accesses to files served by a local kernel NFS server should
+be notified.
+
+In order to make the impact on the file system code as small as possible,
+the problem of hard links to files has been ignored.  So if a file (x)
+exists in two directories (a and b) then a change to the file using the
+name "a/x" should be notified to a program expecting notifications on
+directory "a", but will not be notified to one expecting notifications on
+directory "b".
+
+Also, files that are unlinked, will still cause notifications in the
+last directory that they were linked to.
+
+Configuration
+-------------
+
+Dnotify is controlled via the CONFIG_DNOTIFY configuration option.  When
+disabled, fcntl(fd, F_NOTIFY, ...) will return -EINVAL.
+
+Example
+-------
+See tools/testing/selftests/filesystems/dnotify_test.c for an example.
+
+NOTE
+----
+Beginning with Linux 2.6.13, dnotify has been replaced by inotify.
+See Documentation/filesystems/inotify.txt for more information on it.
diff --git a/Documentation/filesystems/ecryptfs.txt b/Documentation/filesystems/ecryptfs.txt
new file mode 100644
index 000000000..01d8a0835
--- /dev/null
+++ b/Documentation/filesystems/ecryptfs.txt
@@ -0,0 +1,77 @@
+eCryptfs: A stacked cryptographic filesystem for Linux
+
+eCryptfs is free software. Please see the file COPYING for details.
+For documentation, please see the files in the doc/ subdirectory.  For
+building and installation instructions please see the INSTALL file.
+
+Maintainer: Phillip Hellewell
+Lead developer: Michael A. Halcrow <mhalcrow@us.ibm.com>
+Developers: Michael C. Thompson
+            Kent Yoder
+Web Site: http://ecryptfs.sf.net
+
+This software is currently undergoing development. Make sure to
+maintain a backup copy of any data you write into eCryptfs.
+
+eCryptfs requires the userspace tools downloadable from the
+SourceForge site:
+
+http://sourceforge.net/projects/ecryptfs/
+
+Userspace requirements include:
+ - David Howells' userspace keyring headers and libraries (version
+   1.0 or higher), obtainable from
+   http://people.redhat.com/~dhowells/keyutils/
+ - Libgcrypt
+
+
+NOTES
+
+In the beta/experimental releases of eCryptfs, when you upgrade
+eCryptfs, you should copy the files to an unencrypted location and
+then copy the files back into the new eCryptfs mount to migrate the
+files.
+
+
+MOUNT-WIDE PASSPHRASE
+
+Create a new directory into which eCryptfs will write its encrypted
+files (i.e., /root/crypt).  Then, create the mount point directory
+(i.e., /mnt/crypt).  Now it's time to mount eCryptfs:
+
+mount -t ecryptfs /root/crypt /mnt/crypt
+
+You should be prompted for a passphrase and a salt (the salt may be
+blank).
+
+Try writing a new file:
+
+echo "Hello, World" > /mnt/crypt/hello.txt
+
+The operation will complete.  Notice that there is a new file in
+/root/crypt that is at least 12288 bytes in size (depending on your
+host page size).  This is the encrypted underlying file for what you
+just wrote.  To test reading, from start to finish, you need to clear
+the user session keyring:
+
+keyctl clear @u
+
+Then umount /mnt/crypt and mount again per the instructions given
+above.
+
+cat /mnt/crypt/hello.txt
+
+
+NOTES
+
+eCryptfs version 0.1 should only be mounted on (1) empty directories
+or (2) directories containing files only created by eCryptfs. If you
+mount a directory that has pre-existing files not created by eCryptfs,
+then behavior is undefined. Do not run eCryptfs in higher verbosity
+levels unless you are doing so for the sole purpose of debugging or
+development, since secret values will be written out to the system log
+in that case.
+
+
+Mike Halcrow
+mhalcrow@us.ibm.com
diff --git a/Documentation/filesystems/efivarfs.txt b/Documentation/filesystems/efivarfs.txt
new file mode 100644
index 000000000..686a64bba
--- /dev/null
+++ b/Documentation/filesystems/efivarfs.txt
@@ -0,0 +1,23 @@
+
+efivarfs - a (U)EFI variable filesystem
+
+The efivarfs filesystem was created to address the shortcomings of
+using entries in sysfs to maintain EFI variables. The old sysfs EFI
+variables code only supported variables of up to 1024 bytes. This
+limitation existed in version 0.99 of the EFI specification, but was
+removed before any full releases. Since variables can now be larger
+than a single page, sysfs isn't the best interface for this.
+
+Variables can be created, deleted and modified with the efivarfs
+filesystem.
+
+efivarfs is typically mounted like this,
+
+	mount -t efivarfs none /sys/firmware/efi/efivars
+
+Due to the presence of numerous firmware bugs where removing non-standard
+UEFI variables causes the system firmware to fail to POST, efivarfs
+files that are not well-known standardized variables are created
+as immutable files.  This doesn't prevent removal - "chattr -i" will work -
+but it does prevent this kind of failure from being accomplished
+accidentally.
diff --git a/Documentation/filesystems/exofs.txt b/Documentation/filesystems/exofs.txt
new file mode 100644
index 000000000..23583a136
--- /dev/null
+++ b/Documentation/filesystems/exofs.txt
@@ -0,0 +1,185 @@
+===============================================================================
+WHAT IS EXOFS?
+===============================================================================
+
+exofs is a file system that uses an OSD and exports the API of a normal Linux
+file system. Users access exofs like any other local file system, and exofs
+will in turn issue commands to the local OSD initiator.
+
+OSD is a new T10 command set that views storage devices not as a large/flat
+array of sectors but as a container of objects, each having a length, quota,
+time attributes and more. Each object is addressed by a 64bit ID, and is
+contained in a 64bit ID partition. Each object has associated attributes
+attached to it, which are integral part of the object and provide metadata about
+the object. The standard defines some common obligatory attributes, but user
+attributes can be added as needed.
+
+===============================================================================
+ENVIRONMENT
+===============================================================================
+
+To use this file system, you need to have an object store to run it on.  You
+may download a target from:
+http://open-osd.org
+
+See Documentation/scsi/osd.txt for how to setup a working osd environment.
+
+===============================================================================
+USAGE
+===============================================================================
+
+1. Download and compile exofs and open-osd initiator:
+  You need an external Kernel source tree or kernel headers from your
+  distribution. (anything based on 2.6.26 or later).
+
+  a. download open-osd including exofs source using:
+     [parent-directory]$ git clone git://git.open-osd.org/open-osd.git
+
+  b. Build the library module like this:
+     [parent-directory]$ make -C KSRC=$(KER_DIR) open-osd
+
+     This will build both the open-osd initiator as well as the exofs kernel
+     module. Use whatever parameters you compiled your Kernel with and
+     $(KER_DIR) above pointing to the Kernel you compile against. See the file
+     open-osd/top-level-Makefile for an example.
+
+2. Get the OSD initiator and target set up properly, and login to the target.
+  See Documentation/scsi/osd.txt for farther instructions. Also see ./do-osd
+  for example script that does all these steps.
+
+3. Insmod the exofs.ko module:
+   [exofs]$ insmod exofs.ko
+
+4. Make sure the directory where you want to mount exists. If not, create it.
+   (For example, mkdir /mnt/exofs)
+
+5. At first run you will need to invoke the mkfs.exofs application
+
+   As an example, this will create the file system on:
+   /dev/osd0 partition ID 65536
+
+   mkfs.exofs --pid=65536 --format /dev/osd0
+
+   The --format is optional. If not specified, no OSD_FORMAT will be
+   performed and a clean file system will be created in the specified pid,
+   in the available space of the target. (Use --format=size_in_meg to limit
+   the total LUN space available)
+
+   If pid already exists, it will be deleted and a new one will be created in
+   its place. Be careful.
+
+   An exofs lives inside a single OSD partition. You can create multiple exofs
+   filesystems on the same device using multiple pids.
+
+   (run mkfs.exofs without any parameters for usage help message)
+
+6. Mount the file system.
+
+   For example, to mount /dev/osd0, partition ID 0x10000 on /mnt/exofs:
+
+	mount -t exofs -o pid=65536 /dev/osd0 /mnt/exofs/
+
+7. For reference (See do-exofs example script):
+	do-exofs start - an example of how to perform the above steps.
+	do-exofs stop - an example of how to unmount the file system.
+	do-exofs format - an example of how to format and mkfs a new exofs.
+
+8. Extra compilation flags (uncomment in fs/exofs/Kbuild):
+	CONFIG_EXOFS_DEBUG - for debug messages and extra checks.
+
+===============================================================================
+exofs mount options
+===============================================================================
+Similar to any mount command:
+	mount -t exofs -o exofs_options /dev/osdX mount_exofs_directory
+
+Where:
+    -t exofs: specifies the exofs file system
+
+    /dev/osdX: X is a decimal number. /dev/osdX was created after a successful
+               login into an OSD target.
+
+    mount_exofs_directory: The directory to mount the file system on
+
+    exofs specific options: Options are separated by commas (,)
+		pid=<integer> - The partition number to mount/create as
+                                container of the filesystem.
+                                This option is mandatory. integer can be
+                                Hex by pre-pending an 0x to the number.
+		osdname=<id>  - Mount by a device's osdname.
+                                osdname is usually a 36 character uuid of the
+                                form "d2683732-c906-4ee1-9dbd-c10c27bb40df".
+                                It is one of the device's uuid specified in the
+                                mkfs.exofs format command.
+                                If this option is specified then the /dev/osdX
+                                above can be empty and is ignored.
+                to=<integer>  - Timeout in ticks for a single command.
+                                default is (60 * HZ) [for debugging only]
+
+===============================================================================
+DESIGN
+===============================================================================
+
+* The file system control block (AKA on-disk superblock) resides in an object
+  with a special ID (defined in common.h).
+  Information included in the file system control block is used to fill the
+  in-memory superblock structure at mount time. This object is created before
+  the file system is used by mkexofs.c. It contains information such as:
+	- The file system's magic number
+	- The next inode number to be allocated
+
+* Each file resides in its own object and contains the data (and it will be
+  possible to extend the file over multiple objects, though this has not been
+  implemented yet).
+
+* A directory is treated as a file, and essentially contains a list of <file
+  name, inode #> pairs for files that are found in that directory. The object
+  IDs correspond to the files' inode numbers and will be allocated according to
+  a bitmap (stored in a separate object). Now they are allocated using a
+  counter.
+
+* Each file's control block (AKA on-disk inode) is stored in its object's
+  attributes. This applies to both regular files and other types (directories,
+  device files, symlinks, etc.).
+
+* Credentials are generated per object (inode and superblock) when they are
+  created in memory (read from disk or created). The credential works for all
+  operations and is used as long as the object remains in memory.
+
+* Async OSD operations are used whenever possible, but the target may execute
+  them out of order. The operations that concern us are create, delete,
+  readpage, writepage, update_inode, and truncate. The following pairs of
+  operations should execute in the order written, and we need to prevent them
+  from executing in reverse order:
+	- The following are handled with the OBJ_CREATED and OBJ_2BCREATED
+	  flags. OBJ_CREATED is set when we know the object exists on the OSD -
+	  in create's callback function, and when we successfully do a
+	  read_inode.
+	  OBJ_2BCREATED is set in the beginning of the create function, so we
+	  know that we should wait.
+		- create/delete: delete should wait until the object is created
+		  on the OSD.
+		- create/readpage: readpage should be able to return a page
+		  full of zeroes in this case. If there was a write already
+		  en-route (i.e. create, writepage, readpage) then the page
+		  would be locked, and so it would really be the same as
+		  create/writepage.
+		- create/writepage: if writepage is called for a sync write, it
+		  should wait until the object is created on the OSD.
+		  Otherwise, it should just return.
+		- create/truncate: truncate should wait until the object is
+		  created on the OSD.
+		- create/update_inode: update_inode should wait until the
+		  object is created on the OSD.
+	- Handled by VFS locks:
+		- readpage/delete: shouldn't happen because of page lock.
+		- writepage/delete: shouldn't happen because of page lock.
+		- readpage/writepage: shouldn't happen because of page lock.
+
+===============================================================================
+LICENSE/COPYRIGHT
+===============================================================================
+The exofs file system is based on ext2 v0.5b (distributed with the Linux kernel
+version 2.6.10).  All files include the original copyrights, and the license
+is GPL version 2 (only version 2, as is true for the Linux kernel).  The
+Linux kernel can be downloaded from www.kernel.org.
diff --git a/Documentation/filesystems/ext2.txt b/Documentation/filesystems/ext2.txt
new file mode 100644
index 000000000..81c0becab
--- /dev/null
+++ b/Documentation/filesystems/ext2.txt
@@ -0,0 +1,382 @@
+
+The Second Extended Filesystem
+==============================
+
+ext2 was originally released in January 1993.  Written by R\'emy Card,
+Theodore Ts'o and Stephen Tweedie, it was a major rewrite of the
+Extended Filesystem.  It is currently still (April 2001) the predominant
+filesystem in use by Linux.  There are also implementations available
+for NetBSD, FreeBSD, the GNU HURD, Windows 95/98/NT, OS/2 and RISC OS.
+
+Options
+=======
+
+Most defaults are determined by the filesystem superblock, and can be
+set using tune2fs(8). Kernel-determined defaults are indicated by (*).
+
+bsddf			(*)	Makes `df' act like BSD.
+minixdf				Makes `df' act like Minix.
+
+check=none, nocheck	(*)	Don't do extra checking of bitmaps on mount
+				(check=normal and check=strict options removed)
+
+dax				Use direct access (no page cache).  See
+				Documentation/filesystems/dax.txt.
+
+debug				Extra debugging information is sent to the
+				kernel syslog.  Useful for developers.
+
+errors=continue			Keep going on a filesystem error.
+errors=remount-ro		Remount the filesystem read-only on an error.
+errors=panic			Panic and halt the machine if an error occurs.
+
+grpid, bsdgroups		Give objects the same group ID as their parent.
+nogrpid, sysvgroups		New objects have the group ID of their creator.
+
+nouid32				Use 16-bit UIDs and GIDs.
+
+oldalloc			Enable the old block allocator. Orlov should
+				have better performance, we'd like to get some
+				feedback if it's the contrary for you.
+orlov			(*)	Use the Orlov block allocator.
+				(See http://lwn.net/Articles/14633/ and
+				http://lwn.net/Articles/14446/.)
+
+resuid=n			The user ID which may use the reserved blocks.
+resgid=n			The group ID which may use the reserved blocks.
+
+sb=n				Use alternate superblock at this location.
+
+user_xattr			Enable "user." POSIX Extended Attributes
+				(requires CONFIG_EXT2_FS_XATTR).
+nouser_xattr			Don't support "user." extended attributes.
+
+acl				Enable POSIX Access Control Lists support
+				(requires CONFIG_EXT2_FS_POSIX_ACL).
+noacl				Don't support POSIX ACLs.
+
+nobh				Do not attach buffer_heads to file pagecache.
+
+grpquota,noquota,quota,usrquota	Quota options are silently ignored by ext2.
+
+
+Specification
+=============
+
+ext2 shares many properties with traditional Unix filesystems.  It has
+the concepts of blocks, inodes and directories.  It has space in the
+specification for Access Control Lists (ACLs), fragments, undeletion and
+compression though these are not yet implemented (some are available as
+separate patches).  There is also a versioning mechanism to allow new
+features (such as journalling) to be added in a maximally compatible
+manner.
+
+Blocks
+------
+
+The space in the device or file is split up into blocks.  These are
+a fixed size, of 1024, 2048 or 4096 bytes (8192 bytes on Alpha systems),
+which is decided when the filesystem is created.  Smaller blocks mean
+less wasted space per file, but require slightly more accounting overhead,
+and also impose other limits on the size of files and the filesystem.
+
+Block Groups
+------------
+
+Blocks are clustered into block groups in order to reduce fragmentation
+and minimise the amount of head seeking when reading a large amount
+of consecutive data.  Information about each block group is kept in a
+descriptor table stored in the block(s) immediately after the superblock.
+Two blocks near the start of each group are reserved for the block usage
+bitmap and the inode usage bitmap which show which blocks and inodes
+are in use.  Since each bitmap is limited to a single block, this means
+that the maximum size of a block group is 8 times the size of a block.
+
+The block(s) following the bitmaps in each block group are designated
+as the inode table for that block group and the remainder are the data
+blocks.  The block allocation algorithm attempts to allocate data blocks
+in the same block group as the inode which contains them.
+
+The Superblock
+--------------
+
+The superblock contains all the information about the configuration of
+the filing system.  The primary copy of the superblock is stored at an
+offset of 1024 bytes from the start of the device, and it is essential
+to mounting the filesystem.  Since it is so important, backup copies of
+the superblock are stored in block groups throughout the filesystem.
+The first version of ext2 (revision 0) stores a copy at the start of
+every block group, along with backups of the group descriptor block(s).
+Because this can consume a considerable amount of space for large
+filesystems, later revisions can optionally reduce the number of backup
+copies by only putting backups in specific groups (this is the sparse
+superblock feature).  The groups chosen are 0, 1 and powers of 3, 5 and 7.
+
+The information in the superblock contains fields such as the total
+number of inodes and blocks in the filesystem and how many are free,
+how many inodes and blocks are in each block group, when the filesystem
+was mounted (and if it was cleanly unmounted), when it was modified,
+what version of the filesystem it is (see the Revisions section below)
+and which OS created it.
+
+If the filesystem is revision 1 or higher, then there are extra fields,
+such as a volume name, a unique identification number, the inode size,
+and space for optional filesystem features to store configuration info.
+
+All fields in the superblock (as in all other ext2 structures) are stored
+on the disc in little endian format, so a filesystem is portable between
+machines without having to know what machine it was created on.
+
+Inodes
+------
+
+The inode (index node) is a fundamental concept in the ext2 filesystem.
+Each object in the filesystem is represented by an inode.  The inode
+structure contains pointers to the filesystem blocks which contain the
+data held in the object and all of the metadata about an object except
+its name.  The metadata about an object includes the permissions, owner,
+group, flags, size, number of blocks used, access time, change time,
+modification time, deletion time, number of links, fragments, version
+(for NFS) and extended attributes (EAs) and/or Access Control Lists (ACLs).
+
+There are some reserved fields which are currently unused in the inode
+structure and several which are overloaded.  One field is reserved for the
+directory ACL if the inode is a directory and alternately for the top 32
+bits of the file size if the inode is a regular file (allowing file sizes
+larger than 2GB).  The translator field is unused under Linux, but is used
+by the HURD to reference the inode of a program which will be used to
+interpret this object.  Most of the remaining reserved fields have been
+used up for both Linux and the HURD for larger owner and group fields,
+The HURD also has a larger mode field so it uses another of the remaining
+fields to store the extra more bits.
+
+There are pointers to the first 12 blocks which contain the file's data
+in the inode.  There is a pointer to an indirect block (which contains
+pointers to the next set of blocks), a pointer to a doubly-indirect
+block (which contains pointers to indirect blocks) and a pointer to a
+trebly-indirect block (which contains pointers to doubly-indirect blocks).
+
+The flags field contains some ext2-specific flags which aren't catered
+for by the standard chmod flags.  These flags can be listed with lsattr
+and changed with the chattr command, and allow specific filesystem
+behaviour on a per-file basis.  There are flags for secure deletion,
+undeletable, compression, synchronous updates, immutability, append-only,
+dumpable, no-atime, indexed directories, and data-journaling.  Not all
+of these are supported yet.
+
+Directories
+-----------
+
+A directory is a filesystem object and has an inode just like a file.
+It is a specially formatted file containing records which associate
+each name with an inode number.  Later revisions of the filesystem also
+encode the type of the object (file, directory, symlink, device, fifo,
+socket) to avoid the need to check the inode itself for this information
+(support for taking advantage of this feature does not yet exist in
+Glibc 2.2).
+
+The inode allocation code tries to assign inodes which are in the same
+block group as the directory in which they are first created.
+
+The current implementation of ext2 uses a singly-linked list to store
+the filenames in the directory; a pending enhancement uses hashing of the
+filenames to allow lookup without the need to scan the entire directory.
+
+The current implementation never removes empty directory blocks once they
+have been allocated to hold more files.
+
+Special files
+-------------
+
+Symbolic links are also filesystem objects with inodes.  They deserve
+special mention because the data for them is stored within the inode
+itself if the symlink is less than 60 bytes long.  It uses the fields
+which would normally be used to store the pointers to data blocks.
+This is a worthwhile optimisation as it we avoid allocating a full
+block for the symlink, and most symlinks are less than 60 characters long.
+
+Character and block special devices never have data blocks assigned to
+them.  Instead, their device number is stored in the inode, again reusing
+the fields which would be used to point to the data blocks.
+
+Reserved Space
+--------------
+
+In ext2, there is a mechanism for reserving a certain number of blocks
+for a particular user (normally the super-user).  This is intended to
+allow for the system to continue functioning even if non-privileged users
+fill up all the space available to them (this is independent of filesystem
+quotas).  It also keeps the filesystem from filling up entirely which
+helps combat fragmentation.
+
+Filesystem check
+----------------
+
+At boot time, most systems run a consistency check (e2fsck) on their
+filesystems.  The superblock of the ext2 filesystem contains several
+fields which indicate whether fsck should actually run (since checking
+the filesystem at boot can take a long time if it is large).  fsck will
+run if the filesystem was not cleanly unmounted, if the maximum mount
+count has been exceeded or if the maximum time between checks has been
+exceeded.
+
+Feature Compatibility
+---------------------
+
+The compatibility feature mechanism used in ext2 is sophisticated.
+It safely allows features to be added to the filesystem, without
+unnecessarily sacrificing compatibility with older versions of the
+filesystem code.  The feature compatibility mechanism is not supported by
+the original revision 0 (EXT2_GOOD_OLD_REV) of ext2, but was introduced in
+revision 1.  There are three 32-bit fields, one for compatible features
+(COMPAT), one for read-only compatible (RO_COMPAT) features and one for
+incompatible (INCOMPAT) features.
+
+These feature flags have specific meanings for the kernel as follows:
+
+A COMPAT flag indicates that a feature is present in the filesystem,
+but the on-disk format is 100% compatible with older on-disk formats, so
+a kernel which didn't know anything about this feature could read/write
+the filesystem without any chance of corrupting the filesystem (or even
+making it inconsistent).  This is essentially just a flag which says
+"this filesystem has a (hidden) feature" that the kernel or e2fsck may
+want to be aware of (more on e2fsck and feature flags later).  The ext3
+HAS_JOURNAL feature is a COMPAT flag because the ext3 journal is simply
+a regular file with data blocks in it so the kernel does not need to
+take any special notice of it if it doesn't understand ext3 journaling.
+
+An RO_COMPAT flag indicates that the on-disk format is 100% compatible
+with older on-disk formats for reading (i.e. the feature does not change
+the visible on-disk format).  However, an old kernel writing to such a
+filesystem would/could corrupt the filesystem, so this is prevented. The
+most common such feature, SPARSE_SUPER, is an RO_COMPAT feature because
+sparse groups allow file data blocks where superblock/group descriptor
+backups used to live, and ext2_free_blocks() refuses to free these blocks,
+which would leading to inconsistent bitmaps.  An old kernel would also
+get an error if it tried to free a series of blocks which crossed a group
+boundary, but this is a legitimate layout in a SPARSE_SUPER filesystem.
+
+An INCOMPAT flag indicates the on-disk format has changed in some
+way that makes it unreadable by older kernels, or would otherwise
+cause a problem if an old kernel tried to mount it.  FILETYPE is an
+INCOMPAT flag because older kernels would think a filename was longer
+than 256 characters, which would lead to corrupt directory listings.
+The COMPRESSION flag is an obvious INCOMPAT flag - if the kernel
+doesn't understand compression, you would just get garbage back from
+read() instead of it automatically decompressing your data.  The ext3
+RECOVER flag is needed to prevent a kernel which does not understand the
+ext3 journal from mounting the filesystem without replaying the journal.
+
+For e2fsck, it needs to be more strict with the handling of these
+flags than the kernel.  If it doesn't understand ANY of the COMPAT,
+RO_COMPAT, or INCOMPAT flags it will refuse to check the filesystem,
+because it has no way of verifying whether a given feature is valid
+or not.  Allowing e2fsck to succeed on a filesystem with an unknown
+feature is a false sense of security for the user.  Refusing to check
+a filesystem with unknown features is a good incentive for the user to
+update to the latest e2fsck.  This also means that anyone adding feature
+flags to ext2 also needs to update e2fsck to verify these features.
+
+Metadata
+--------
+
+It is frequently claimed that the ext2 implementation of writing
+asynchronous metadata is faster than the ffs synchronous metadata
+scheme but less reliable.  Both methods are equally resolvable by their
+respective fsck programs.
+
+If you're exceptionally paranoid, there are 3 ways of making metadata
+writes synchronous on ext2:
+
+per-file if you have the program source: use the O_SYNC flag to open()
+per-file if you don't have the source: use "chattr +S" on the file
+per-filesystem: add the "sync" option to mount (or in /etc/fstab)
+
+the first and last are not ext2 specific but do force the metadata to
+be written synchronously.  See also Journaling below.
+
+Limitations
+-----------
+
+There are various limits imposed by the on-disk layout of ext2.  Other
+limits are imposed by the current implementation of the kernel code.
+Many of the limits are determined at the time the filesystem is first
+created, and depend upon the block size chosen.  The ratio of inodes to
+data blocks is fixed at filesystem creation time, so the only way to
+increase the number of inodes is to increase the size of the filesystem.
+No tools currently exist which can change the ratio of inodes to blocks.
+
+Most of these limits could be overcome with slight changes in the on-disk
+format and using a compatibility flag to signal the format change (at
+the expense of some compatibility).
+
+Filesystem block size:     1kB        2kB        4kB        8kB
+
+File size limit:          16GB      256GB     2048GB     2048GB
+Filesystem size limit:  2047GB     8192GB    16384GB    32768GB
+
+There is a 2.4 kernel limit of 2048GB for a single block device, so no
+filesystem larger than that can be created at this time.  There is also
+an upper limit on the block size imposed by the page size of the kernel,
+so 8kB blocks are only allowed on Alpha systems (and other architectures
+which support larger pages).
+
+There is an upper limit of 32000 subdirectories in a single directory.
+
+There is a "soft" upper limit of about 10-15k files in a single directory
+with the current linear linked-list directory implementation.  This limit
+stems from performance problems when creating and deleting (and also
+finding) files in such large directories.  Using a hashed directory index
+(under development) allows 100k-1M+ files in a single directory without
+performance problems (although RAM size becomes an issue at this point).
+
+The (meaningless) absolute upper limit of files in a single directory
+(imposed by the file size, the realistic limit is obviously much less)
+is over 130 trillion files.  It would be higher except there are not
+enough 4-character names to make up unique directory entries, so they
+have to be 8 character filenames, even then we are fairly close to
+running out of unique filenames.
+
+Journaling
+----------
+
+A journaling extension to the ext2 code has been developed by Stephen
+Tweedie.  It avoids the risks of metadata corruption and the need to
+wait for e2fsck to complete after a crash, without requiring a change
+to the on-disk ext2 layout.  In a nutshell, the journal is a regular
+file which stores whole metadata (and optionally data) blocks that have
+been modified, prior to writing them into the filesystem.  This means
+it is possible to add a journal to an existing ext2 filesystem without
+the need for data conversion.
+
+When changes to the filesystem (e.g. a file is renamed) they are stored in
+a transaction in the journal and can either be complete or incomplete at
+the time of a crash.  If a transaction is complete at the time of a crash
+(or in the normal case where the system does not crash), then any blocks
+in that transaction are guaranteed to represent a valid filesystem state,
+and are copied into the filesystem.  If a transaction is incomplete at
+the time of the crash, then there is no guarantee of consistency for
+the blocks in that transaction so they are discarded (which means any
+filesystem changes they represent are also lost).
+Check Documentation/filesystems/ext4.txt if you want to read more about
+ext4 and journaling.
+
+References
+==========
+
+The kernel source	file:/usr/src/linux/fs/ext2/
+e2fsprogs (e2fsck)	http://e2fsprogs.sourceforge.net/
+Design & Implementation	http://e2fsprogs.sourceforge.net/ext2intro.html
+Journaling (ext3)	ftp://ftp.uk.linux.org/pub/linux/sct/fs/jfs/
+Filesystem Resizing	http://ext2resize.sourceforge.net/
+Compression (*)		http://e2compr.sourceforge.net/
+
+Implementations for:
+Windows 95/98/NT/2000	http://www.chrysocome.net/explore2fs
+Windows 95 (*)		http://www.yipton.net/content.html#FSDEXT2
+DOS client (*)		ftp://metalab.unc.edu/pub/Linux/system/filesystems/ext2/
+OS/2 (+)		ftp://metalab.unc.edu/pub/Linux/system/filesystems/ext2/
+RISC OS client		http://www.esw-heim.tu-clausthal.de/~marco/smorbrod/IscaFS/
+
+(*) no longer actively developed/supported (as of Apr 2001)
+(+) no longer actively developed/supported (as of Mar 2009)
diff --git a/Documentation/filesystems/ext3.txt b/Documentation/filesystems/ext3.txt
new file mode 100644
index 000000000..58758fbef
--- /dev/null
+++ b/Documentation/filesystems/ext3.txt
@@ -0,0 +1,12 @@
+
+Ext3 Filesystem
+===============
+
+Ext3 was originally released in September 1999. Written by Stephen Tweedie
+for the 2.2 branch, and ported to 2.4 kernels by Peter Braam, Andreas Dilger,
+Andrew Morton, Alexander Viro, Ted Ts'o and Stephen Tweedie.
+
+Ext3 is the ext2 filesystem enhanced with journalling capabilities. The
+filesystem is a subset of ext4 filesystem so use ext4 driver for accessing
+ext3 filesystems.
+
diff --git a/Documentation/filesystems/ext4/ext4.rst b/Documentation/filesystems/ext4/ext4.rst
new file mode 100644
index 000000000..9d4368d59
--- /dev/null
+++ b/Documentation/filesystems/ext4/ext4.rst
@@ -0,0 +1,613 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+========================
+General Information
+========================
+
+Ext4 is an advanced level of the ext3 filesystem which incorporates
+scalability and reliability enhancements for supporting large filesystems
+(64 bit) in keeping with increasing disk capacities and state-of-the-art
+feature requirements.
+
+Mailing list:	linux-ext4@vger.kernel.org
+Web site:	http://ext4.wiki.kernel.org
+
+
+Quick usage instructions
+========================
+
+Note: More extensive information for getting started with ext4 can be
+found at the ext4 wiki site at the URL:
+http://ext4.wiki.kernel.org/index.php/Ext4_Howto
+
+  - The latest version of e2fsprogs can be found at:
+
+    https://www.kernel.org/pub/linux/kernel/people/tytso/e2fsprogs/
+
+	or
+
+    http://sourceforge.net/project/showfiles.php?group_id=2406
+
+	or grab the latest git repository from:
+
+   https://git.kernel.org/pub/scm/fs/ext2/e2fsprogs.git
+
+  - Create a new filesystem using the ext4 filesystem type:
+
+        # mke2fs -t ext4 /dev/hda1
+
+    Or to configure an existing ext3 filesystem to support extents:
+
+	# tune2fs -O extents /dev/hda1
+
+    If the filesystem was created with 128 byte inodes, it can be
+    converted to use 256 byte for greater efficiency via:
+
+        # tune2fs -I 256 /dev/hda1
+
+  - Mounting:
+
+	# mount -t ext4 /dev/hda1 /wherever
+
+  - When comparing performance with other filesystems, it's always
+    important to try multiple workloads; very often a subtle change in a
+    workload parameter can completely change the ranking of which
+    filesystems do well compared to others.  When comparing versus ext3,
+    note that ext4 enables write barriers by default, while ext3 does
+    not enable write barriers by default.  So it is useful to use
+    explicitly specify whether barriers are enabled or not when via the
+    '-o barriers=[0|1]' mount option for both ext3 and ext4 filesystems
+    for a fair comparison.  When tuning ext3 for best benchmark numbers,
+    it is often worthwhile to try changing the data journaling mode; '-o
+    data=writeback' can be faster for some workloads.  (Note however that
+    running mounted with data=writeback can potentially leave stale data
+    exposed in recently written files in case of an unclean shutdown,
+    which could be a security exposure in some situations.)  Configuring
+    the filesystem with a large journal can also be helpful for
+    metadata-intensive workloads.
+
+Features
+========
+
+Currently Available
+-------------------
+
+* ability to use filesystems > 16TB (e2fsprogs support not available yet)
+* extent format reduces metadata overhead (RAM, IO for access, transactions)
+* extent format more robust in face of on-disk corruption due to magics,
+* internal redundancy in tree
+* improved file allocation (multi-block alloc)
+* lift 32000 subdirectory limit imposed by i_links_count[1]
+* nsec timestamps for mtime, atime, ctime, create time
+* inode version field on disk (NFSv4, Lustre)
+* reduced e2fsck time via uninit_bg feature
+* journal checksumming for robustness, performance
+* persistent file preallocation (e.g for streaming media, databases)
+* ability to pack bitmaps and inode tables into larger virtual groups via the
+  flex_bg feature
+* large file support
+* inode allocation using large virtual block groups via flex_bg
+* delayed allocation
+* large block (up to pagesize) support
+* efficient new ordered mode in JBD2 and ext4 (avoid using buffer head to force
+  the ordering)
+
+[1] Filesystems with a block size of 1k may see a limit imposed by the
+directory hash tree having a maximum depth of two.
+
+Options
+=======
+
+When mounting an ext4 filesystem, the following option are accepted:
+(*) == default
+
+======================= =======================================================
+Mount Option            Description
+======================= =======================================================
+ro                   	Mount filesystem read only. Note that ext4 will
+                     	replay the journal (and thus write to the
+                     	partition) even when mounted "read only". The
+                     	mount options "ro,noload" can be used to prevent
+		     	writes to the filesystem.
+
+journal_checksum	Enable checksumming of the journal transactions.
+			This will allow the recovery code in e2fsck and the
+			kernel to detect corruption in the kernel.  It is a
+			compatible change and will be ignored by older kernels.
+
+journal_async_commit	Commit block can be written to disk without waiting
+			for descriptor blocks. If enabled older kernels cannot
+			mount the device. This will enable 'journal_checksum'
+			internally.
+
+journal_path=path
+journal_dev=devnum	When the external journal device's major/minor numbers
+			have changed, these options allow the user to specify
+			the new journal location.  The journal device is
+			identified through either its new major/minor numbers
+			encoded in devnum, or via a path to the device.
+
+norecovery		Don't load the journal on mounting.  Note that
+noload			if the filesystem was not unmounted cleanly,
+                     	skipping the journal replay will lead to the
+                     	filesystem containing inconsistencies that can
+                     	lead to any number of problems.
+
+data=journal		All data are committed into the journal prior to being
+			written into the main file system.  Enabling
+			this mode will disable delayed allocation and
+			O_DIRECT support.
+
+data=ordered	(*)	All data are forced directly out to the main file
+			system prior to its metadata being committed to the
+			journal.
+
+data=writeback		Data ordering is not preserved, data may be written
+			into the main file system after its metadata has been
+			committed to the journal.
+
+commit=nrsec	(*)	Ext4 can be told to sync all its data and metadata
+			every 'nrsec' seconds. The default value is 5 seconds.
+			This means that if you lose your power, you will lose
+			as much as the latest 5 seconds of work (your
+			filesystem will not be damaged though, thanks to the
+			journaling).  This default value (or any low value)
+			will hurt performance, but it's good for data-safety.
+			Setting it to 0 will have the same effect as leaving
+			it at the default (5 seconds).
+			Setting it to very large values will improve
+			performance.
+
+barrier=<0|1(*)>	This enables/disables the use of write barriers in
+barrier(*)		the jbd code.  barrier=0 disables, barrier=1 enables.
+nobarrier		This also requires an IO stack which can support
+			barriers, and if jbd gets an error on a barrier
+			write, it will disable again with a warning.
+			Write barriers enforce proper on-disk ordering
+			of journal commits, making volatile disk write caches
+			safe to use, at some performance penalty.  If
+			your disks are battery-backed in one way or another,
+			disabling barriers may safely improve performance.
+			The mount options "barrier" and "nobarrier" can
+			also be used to enable or disable barriers, for
+			consistency with other ext4 mount options.
+
+inode_readahead_blks=n	This tuning parameter controls the maximum
+			number of inode table blocks that ext4's inode
+			table readahead algorithm will pre-read into
+			the buffer cache.  The default value is 32 blocks.
+
+nouser_xattr		Disables Extended User Attributes.  See the
+			attr(5) manual page for more information about
+			extended attributes.
+
+noacl			This option disables POSIX Access Control List
+			support. If ACL support is enabled in the kernel
+			configuration (CONFIG_EXT4_FS_POSIX_ACL), ACL is
+			enabled by default on mount. See the acl(5) manual
+			page for more information about acl.
+
+bsddf		(*)	Make 'df' act like BSD.
+minixdf			Make 'df' act like Minix.
+
+debug			Extra debugging information is sent to syslog.
+
+abort			Simulate the effects of calling ext4_abort() for
+			debugging purposes.  This is normally used while
+			remounting a filesystem which is already mounted.
+
+errors=remount-ro	Remount the filesystem read-only on an error.
+errors=continue		Keep going on a filesystem error.
+errors=panic		Panic and halt the machine if an error occurs.
+                        (These mount options override the errors behavior
+                        specified in the superblock, which can be configured
+                        using tune2fs)
+
+data_err=ignore(*)	Just print an error message if an error occurs
+			in a file data buffer in ordered mode.
+data_err=abort		Abort the journal if an error occurs in a file
+			data buffer in ordered mode.
+
+grpid			New objects have the group ID of their parent.
+bsdgroups
+
+nogrpid		(*)	New objects have the group ID of their creator.
+sysvgroups
+
+resgid=n		The group ID which may use the reserved blocks.
+
+resuid=n		The user ID which may use the reserved blocks.
+
+sb=n			Use alternate superblock at this location.
+
+quota			These options are ignored by the filesystem. They
+noquota			are used only by quota tools to recognize volumes
+grpquota		where quota should be turned on. See documentation
+usrquota		in the quota-tools package for more details
+			(http://sourceforge.net/projects/linuxquota).
+
+jqfmt=<quota type>	These options tell filesystem details about quota
+usrjquota=<file>	so that quota information can be properly updated
+grpjquota=<file>	during journal replay. They replace the above
+			quota options. See documentation in the quota-tools
+			package for more details
+			(http://sourceforge.net/projects/linuxquota).
+
+stripe=n		Number of filesystem blocks that mballoc will try
+			to use for allocation size and alignment. For RAID5/6
+			systems this should be the number of data
+			disks *  RAID chunk size in file system blocks.
+
+delalloc	(*)	Defer block allocation until just before ext4
+			writes out the block(s) in question.  This
+			allows ext4 to better allocation decisions
+			more efficiently.
+nodelalloc		Disable delayed allocation.  Blocks are allocated
+			when the data is copied from userspace to the
+			page cache, either via the write(2) system call
+			or when an mmap'ed page which was previously
+			unallocated is written for the first time.
+
+max_batch_time=usec	Maximum amount of time ext4 should wait for
+			additional filesystem operations to be batch
+			together with a synchronous write operation.
+			Since a synchronous write operation is going to
+			force a commit and then a wait for the I/O
+			complete, it doesn't cost much, and can be a
+			huge throughput win, we wait for a small amount
+			of time to see if any other transactions can
+			piggyback on the synchronous write.   The
+			algorithm used is designed to automatically tune
+			for the speed of the disk, by measuring the
+			amount of time (on average) that it takes to
+			finish committing a transaction.  Call this time
+			the "commit time".  If the time that the
+			transaction has been running is less than the
+			commit time, ext4 will try sleeping for the
+			commit time to see if other operations will join
+			the transaction.   The commit time is capped by
+			the max_batch_time, which defaults to 15000us
+			(15ms).   This optimization can be turned off
+			entirely by setting max_batch_time to 0.
+
+min_batch_time=usec	This parameter sets the commit time (as
+			described above) to be at least min_batch_time.
+			It defaults to zero microseconds.  Increasing
+			this parameter may improve the throughput of
+			multi-threaded, synchronous workloads on very
+			fast disks, at the cost of increasing latency.
+
+journal_ioprio=prio	The I/O priority (from 0 to 7, where 0 is the
+			highest priority) which should be used for I/O
+			operations submitted by kjournald2 during a
+			commit operation.  This defaults to 3, which is
+			a slightly higher priority than the default I/O
+			priority.
+
+auto_da_alloc(*)	Many broken applications don't use fsync() when 
+noauto_da_alloc		replacing existing files via patterns such as
+			fd = open("foo.new")/write(fd,..)/close(fd)/
+			rename("foo.new", "foo"), or worse yet,
+			fd = open("foo", O_TRUNC)/write(fd,..)/close(fd).
+			If auto_da_alloc is enabled, ext4 will detect
+			the replace-via-rename and replace-via-truncate
+			patterns and force that any delayed allocation
+			blocks are allocated such that at the next
+			journal commit, in the default data=ordered
+			mode, the data blocks of the new file are forced
+			to disk before the rename() operation is
+			committed.  This provides roughly the same level
+			of guarantees as ext3, and avoids the
+			"zero-length" problem that can happen when a
+			system crashes before the delayed allocation
+			blocks are forced to disk.
+
+noinit_itable		Do not initialize any uninitialized inode table
+			blocks in the background.  This feature may be
+			used by installation CD's so that the install
+			process can complete as quickly as possible; the
+			inode table initialization process would then be
+			deferred until the next time the  file system
+			is unmounted.
+
+init_itable=n		The lazy itable init code will wait n times the
+			number of milliseconds it took to zero out the
+			previous block group's inode table.  This
+			minimizes the impact on the system performance
+			while file system's inode table is being initialized.
+
+discard			Controls whether ext4 should issue discard/TRIM
+nodiscard(*)		commands to the underlying block device when
+			blocks are freed.  This is useful for SSD devices
+			and sparse/thinly-provisioned LUNs, but it is off
+			by default until sufficient testing has been done.
+
+nouid32			Disables 32-bit UIDs and GIDs.  This is for
+			interoperability  with  older kernels which only
+			store and expect 16-bit values.
+
+block_validity(*)	These options enable or disable the in-kernel
+noblock_validity	facility for tracking filesystem metadata blocks
+			within internal data structures.  This allows multi-
+			block allocator and other routines to notice
+			bugs or corrupted allocation bitmaps which cause
+			blocks to be allocated which overlap with
+			filesystem metadata blocks.
+
+dioread_lock		Controls whether or not ext4 should use the DIO read
+dioread_nolock		locking. If the dioread_nolock option is specified
+			ext4 will allocate uninitialized extent before buffer
+			write and convert the extent to initialized after IO
+			completes. This approach allows ext4 code to avoid
+			using inode mutex, which improves scalability on high
+			speed storages. However this does not work with
+			data journaling and dioread_nolock option will be
+			ignored with kernel warning. Note that dioread_nolock
+			code path is only used for extent-based files.
+			Because of the restrictions this options comprises
+			it is off by default (e.g. dioread_lock).
+
+max_dir_size_kb=n	This limits the size of directories so that any
+			attempt to expand them beyond the specified
+			limit in kilobytes will cause an ENOSPC error.
+			This is useful in memory constrained
+			environments, where a very large directory can
+			cause severe performance problems or even
+			provoke the Out Of Memory killer.  (For example,
+			if there is only 512mb memory available, a 176mb
+			directory may seriously cramp the system's style.)
+
+i_version		Enable 64-bit inode version support. This option is
+			off by default.
+
+dax			Use direct access (no page cache).  See
+			Documentation/filesystems/dax.txt.  Note that
+			this option is incompatible with data=journal.
+======================= =======================================================
+
+Data Mode
+=========
+There are 3 different data modes:
+
+* writeback mode
+
+  In data=writeback mode, ext4 does not journal data at all.  This mode provides
+  a similar level of journaling as that of XFS, JFS, and ReiserFS in its default
+  mode - metadata journaling.  A crash+recovery can cause incorrect data to
+  appear in files which were written shortly before the crash.  This mode will
+  typically provide the best ext4 performance.
+
+* ordered mode
+
+  In data=ordered mode, ext4 only officially journals metadata, but it logically
+  groups metadata information related to data changes with the data blocks into
+  a single unit called a transaction.  When it's time to write the new metadata
+  out to disk, the associated data blocks are written first.  In general, this
+  mode performs slightly slower than writeback but significantly faster than
+  journal mode.
+
+* journal mode
+
+  data=journal mode provides full data and metadata journaling.  All new data is
+  written to the journal first, and then to its final location.  In the event of
+  a crash, the journal can be replayed, bringing both data and metadata into a
+  consistent state.  This mode is the slowest except when data needs to be read
+  from and written to disk at the same time where it outperforms all others
+  modes.  Enabling this mode will disable delayed allocation and O_DIRECT
+  support.
+
+/proc entries
+=============
+
+Information about mounted ext4 file systems can be found in
+/proc/fs/ext4.  Each mounted filesystem will have a directory in
+/proc/fs/ext4 based on its device name (i.e., /proc/fs/ext4/hdc or
+/proc/fs/ext4/dm-0).   The files in each per-device directory are shown
+in table below.
+
+Files in /proc/fs/ext4/<devname>
+
+================ =======
+ File            Content
+================ =======
+ mb_groups       details of multiblock allocator buddy cache of free blocks
+================ =======
+
+/sys entries
+============
+
+Information about mounted ext4 file systems can be found in
+/sys/fs/ext4.  Each mounted filesystem will have a directory in
+/sys/fs/ext4 based on its device name (i.e., /sys/fs/ext4/hdc or
+/sys/fs/ext4/dm-0).   The files in each per-device directory are shown
+in table below.
+
+Files in /sys/fs/ext4/<devname>:
+
+(see also Documentation/ABI/testing/sysfs-fs-ext4)
+
+============================= =================================================
+File                          Content
+============================= =================================================
+ delayed_allocation_blocks    This file is read-only and shows the number of
+                              blocks that are dirty in the page cache, but
+                              which do not have their location in the
+                              filesystem allocated yet.
+
+inode_goal                    Tuning parameter which (if non-zero) controls
+                              the goal inode used by the inode allocator in
+                              preference to all other allocation heuristics.
+                              This is intended for debugging use only, and
+                              should be 0 on production systems.
+
+inode_readahead_blks          Tuning parameter which controls the maximum
+                              number of inode table blocks that ext4's inode
+                              table readahead algorithm will pre-read into
+                              the buffer cache
+
+lifetime_write_kbytes         This file is read-only and shows the number of
+                              kilobytes of data that have been written to this
+                              filesystem since it was created.
+
+ max_writeback_mb_bump        The maximum number of megabytes the writeback
+                              code will try to write out before move on to
+                              another inode.
+
+ mb_group_prealloc            The multiblock allocator will round up allocation
+                              requests to a multiple of this tuning parameter if
+                              the stripe size is not set in the ext4 superblock
+
+ mb_max_to_scan               The maximum number of extents the multiblock
+                              allocator will search to find the best extent
+
+ mb_min_to_scan               The minimum number of extents the multiblock
+                              allocator will search to find the best extent
+
+ mb_order2_req                Tuning parameter which controls the minimum size
+                              for requests (as a power of 2) where the buddy
+                              cache is used
+
+ mb_stats                     Controls whether the multiblock allocator should
+                              collect statistics, which are shown during the
+                              unmount. 1 means to collect statistics, 0 means
+                              not to collect statistics
+
+ mb_stream_req                Files which have fewer blocks than this tunable
+                              parameter will have their blocks allocated out
+                              of a block group specific preallocation pool, so
+                              that small files are packed closely together.
+                              Each large file will have its blocks allocated
+                              out of its own unique preallocation pool.
+
+ session_write_kbytes         This file is read-only and shows the number of
+                              kilobytes of data that have been written to this
+                              filesystem since it was mounted.
+
+ reserved_clusters            This is RW file and contains number of reserved
+                              clusters in the file system which will be used
+                              in the specific situations to avoid costly
+                              zeroout, unexpected ENOSPC, or possible data
+                              loss. The default is 2% or 4096 clusters,
+                              whichever is smaller and this can be changed
+                              however it can never exceed number of clusters
+                              in the file system. If there is not enough space
+                              for the reserved space when mounting the file
+                              mount will _not_ fail.
+============================= =================================================
+
+Ioctls
+======
+
+There is some Ext4 specific functionality which can be accessed by applications
+through the system call interfaces. The list of all Ext4 specific ioctls are
+shown in the table below.
+
+Table of Ext4 specific ioctls
+
+============================= =================================================
+Ioctl			      Description
+============================= =================================================
+ EXT4_IOC_GETFLAGS	      Get additional attributes associated with inode.
+			      The ioctl argument is an integer bitfield, with
+			      bit values described in ext4.h. This ioctl is an
+			      alias for FS_IOC_GETFLAGS.
+
+ EXT4_IOC_SETFLAGS	      Set additional attributes associated with inode.
+			      The ioctl argument is an integer bitfield, with
+			      bit values described in ext4.h. This ioctl is an
+			      alias for FS_IOC_SETFLAGS.
+
+ EXT4_IOC_GETVERSION
+ EXT4_IOC_GETVERSION_OLD
+			      Get the inode i_generation number stored for
+			      each inode. The i_generation number is normally
+			      changed only when new inode is created and it is
+			      particularly useful for network filesystems. The
+			      '_OLD' version of this ioctl is an alias for
+			      FS_IOC_GETVERSION.
+
+ EXT4_IOC_SETVERSION
+ EXT4_IOC_SETVERSION_OLD
+			      Set the inode i_generation number stored for
+			      each inode. The '_OLD' version of this ioctl
+			      is an alias for FS_IOC_SETVERSION.
+
+ EXT4_IOC_GROUP_EXTEND	      This ioctl has the same purpose as the resize
+			      mount option. It allows to resize filesystem
+			      to the end of the last existing block group,
+			      further resize has to be done with resize2fs,
+			      either online, or offline. The argument points
+			      to the unsigned logn number representing the
+			      filesystem new block count.
+
+ EXT4_IOC_MOVE_EXT	      Move the block extents from orig_fd (the one
+			      this ioctl is pointing to) to the donor_fd (the
+			      one specified in move_extent structure passed
+			      as an argument to this ioctl). Then, exchange
+			      inode metadata between orig_fd and donor_fd.
+			      This is especially useful for online
+			      defragmentation, because the allocator has the
+			      opportunity to allocate moved blocks better,
+			      ideally into one contiguous extent.
+
+ EXT4_IOC_GROUP_ADD	      Add a new group descriptor to an existing or
+			      new group descriptor block. The new group
+			      descriptor is described by ext4_new_group_input
+			      structure, which is passed as an argument to
+			      this ioctl. This is especially useful in
+			      conjunction with EXT4_IOC_GROUP_EXTEND,
+			      which allows online resize of the filesystem
+			      to the end of the last existing block group.
+			      Those two ioctls combined is used in userspace
+			      online resize tool (e.g. resize2fs).
+
+ EXT4_IOC_MIGRATE	      This ioctl operates on the filesystem itself.
+			      It converts (migrates) ext3 indirect block mapped
+			      inode to ext4 extent mapped inode by walking
+			      through indirect block mapping of the original
+			      inode and converting contiguous block ranges
+			      into ext4 extents of the temporary inode. Then,
+			      inodes are swapped. This ioctl might help, when
+			      migrating from ext3 to ext4 filesystem, however
+			      suggestion is to create fresh ext4 filesystem
+			      and copy data from the backup. Note, that
+			      filesystem has to support extents for this ioctl
+			      to work.
+
+ EXT4_IOC_ALLOC_DA_BLKS	      Force all of the delay allocated blocks to be
+			      allocated to preserve application-expected ext3
+			      behaviour. Note that this will also start
+			      triggering a write of the data blocks, but this
+			      behaviour may change in the future as it is
+			      not necessary and has been done this way only
+			      for sake of simplicity.
+
+ EXT4_IOC_RESIZE_FS	      Resize the filesystem to a new size.  The number
+			      of blocks of resized filesystem is passed in via
+			      64 bit integer argument.  The kernel allocates
+			      bitmaps and inode table, the userspace tool thus
+			      just passes the new number of blocks.
+
+ EXT4_IOC_SWAP_BOOT	      Swap i_blocks and associated attributes
+			      (like i_blocks, i_size, i_flags, ...) from
+			      the specified inode with inode
+			      EXT4_BOOT_LOADER_INO (#5). This is typically
+			      used to store a boot loader in a secure part of
+			      the filesystem, where it can't be changed by a
+			      normal user by accident.
+			      The data blocks of the previous boot loader
+			      will be associated with the given inode.
+============================= =================================================
+
+References
+==========
+
+kernel source:	<file:fs/ext4/>
+		<file:fs/jbd2/>
+
+programs:	http://e2fsprogs.sourceforge.net/
+
+useful links:	http://fedoraproject.org/wiki/ext3-devel
+		http://www.bullopensource.org/ext4/
+		http://ext4.wiki.kernel.org/index.php/Main_Page
+		http://fedoraproject.org/wiki/Features/Ext4
diff --git a/Documentation/filesystems/ext4/index.rst b/Documentation/filesystems/ext4/index.rst
new file mode 100644
index 000000000..711216055
--- /dev/null
+++ b/Documentation/filesystems/ext4/index.rst
@@ -0,0 +1,17 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+===============
+ext4 Filesystem
+===============
+
+General usage and on-disk artifacts writen by ext4.  More documentation may
+be ported from the wiki as time permits.  This should be considered the
+canonical source of information as the details here have been reviewed by
+the ext4 community.
+
+.. toctree::
+   :maxdepth: 5
+   :numbered:
+
+   ext4
+   ondisk/index
diff --git a/Documentation/filesystems/ext4/ondisk/about.rst b/Documentation/filesystems/ext4/ondisk/about.rst
new file mode 100644
index 000000000..0aadba052
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/about.rst
@@ -0,0 +1,44 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+About this Book
+===============
+
+This document attempts to describe the on-disk format for ext4
+filesystems. The same general ideas should apply to ext2/3 filesystems
+as well, though they do not support all the features that ext4 supports,
+and the fields will be shorter.
+
+**NOTE**: This is a work in progress, based on notes that the author
+(djwong) made while picking apart a filesystem by hand. The data
+structure definitions should be current as of Linux 4.18 and
+e2fsprogs-1.44. All comments and corrections are welcome, since there is
+undoubtedly plenty of lore that might not be reflected in freshly
+created demonstration filesystems.
+
+License
+-------
+This book is licensed under the terms of the GNU Public License, v2.
+
+Terminology
+-----------
+
+ext4 divides a storage device into an array of logical blocks both to
+reduce bookkeeping overhead and to increase throughput by forcing larger
+transfer sizes. Generally, the block size will be 4KiB (the same size as
+pages on x86 and the block layer's default block size), though the
+actual size is calculated as 2 ^ (10 + ``sb.s_log_block_size``) bytes.
+Throughout this document, disk locations are given in terms of these
+logical blocks, not raw LBAs, and not 1024-byte blocks. For the sake of
+convenience, the logical block size will be referred to as
+``$block_size`` throughout the rest of the document.
+
+When referenced in ``preformatted text`` blocks, ``sb`` refers to fields
+in the super block, and ``inode`` refers to fields in an inode table
+entry.
+
+Other References
+----------------
+
+Also see http://www.nongnu.org/ext2-doc/ for quite a collection of
+information about ext2/3. Here's another old reference:
+http://wiki.osdev.org/Ext2
diff --git a/Documentation/filesystems/ext4/ondisk/allocators.rst b/Documentation/filesystems/ext4/ondisk/allocators.rst
new file mode 100644
index 000000000..7aa85152a
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/allocators.rst
@@ -0,0 +1,56 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Block and Inode Allocation Policy
+---------------------------------
+
+ext4 recognizes (better than ext3, anyway) that data locality is
+generally a desirably quality of a filesystem. On a spinning disk,
+keeping related blocks near each other reduces the amount of movement
+that the head actuator and disk must perform to access a data block,
+thus speeding up disk IO. On an SSD there of course are no moving parts,
+but locality can increase the size of each transfer request while
+reducing the total number of requests. This locality may also have the
+effect of concentrating writes on a single erase block, which can speed
+up file rewrites significantly. Therefore, it is useful to reduce
+fragmentation whenever possible.
+
+The first tool that ext4 uses to combat fragmentation is the multi-block
+allocator. When a file is first created, the block allocator
+speculatively allocates 8KiB of disk space to the file on the assumption
+that the space will get written soon. When the file is closed, the
+unused speculative allocations are of course freed, but if the
+speculation is correct (typically the case for full writes of small
+files) then the file data gets written out in a single multi-block
+extent. A second related trick that ext4 uses is delayed allocation.
+Under this scheme, when a file needs more blocks to absorb file writes,
+the filesystem defers deciding the exact placement on the disk until all
+the dirty buffers are being written out to disk. By not committing to a
+particular placement until it's absolutely necessary (the commit timeout
+is hit, or sync() is called, or the kernel runs out of memory), the hope
+is that the filesystem can make better location decisions.
+
+The third trick that ext4 (and ext3) uses is that it tries to keep a
+file's data blocks in the same block group as its inode. This cuts down
+on the seek penalty when the filesystem first has to read a file's inode
+to learn where the file's data blocks live and then seek over to the
+file's data blocks to begin I/O operations.
+
+The fourth trick is that all the inodes in a directory are placed in the
+same block group as the directory, when feasible. The working assumption
+here is that all the files in a directory might be related, therefore it
+is useful to try to keep them all together.
+
+The fifth trick is that the disk volume is cut up into 128MB block
+groups; these mini-containers are used as outlined above to try to
+maintain data locality. However, there is a deliberate quirk -- when a
+directory is created in the root directory, the inode allocator scans
+the block groups and puts that directory into the least heavily loaded
+block group that it can find. This encourages directories to spread out
+over a disk; as the top-level directory/file blobs fill up one block
+group, the allocators simply move on to the next block group. Allegedly
+this scheme evens out the loading on the block groups, though the author
+suspects that the directories which are so unlucky as to land towards
+the end of a spinning drive get a raw deal performance-wise.
+
+Of course if all of these mechanisms fail, one can always use e4defrag
+to defragment files.
diff --git a/Documentation/filesystems/ext4/ondisk/attributes.rst b/Documentation/filesystems/ext4/ondisk/attributes.rst
new file mode 100644
index 000000000..0b01b67b8
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/attributes.rst
@@ -0,0 +1,191 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Extended Attributes
+-------------------
+
+Extended attributes (xattrs) are typically stored in a separate data
+block on the disk and referenced from inodes via ``inode.i_file_acl*``.
+The first use of extended attributes seems to have been for storing file
+ACLs and other security data (selinux). With the ``user_xattr`` mount
+option it is possible for users to store extended attributes so long as
+all attribute names begin with “user”; this restriction seems to have
+disappeared as of Linux 3.0.
+
+There are two places where extended attributes can be found. The first
+place is between the end of each inode entry and the beginning of the
+next inode entry. For example, if inode.i\_extra\_isize = 28 and
+sb.inode\_size = 256, then there are 256 - (128 + 28) = 100 bytes
+available for in-inode extended attribute storage. The second place
+where extended attributes can be found is in the block pointed to by
+``inode.i_file_acl``. As of Linux 3.11, it is not possible for this
+block to contain a pointer to a second extended attribute block (or even
+the remaining blocks of a cluster). In theory it is possible for each
+attribute's value to be stored in a separate data block, though as of
+Linux 3.11 the code does not permit this.
+
+Keys are generally assumed to be ASCIIZ strings, whereas values can be
+strings or binary data.
+
+Extended attributes, when stored after the inode, have a header
+``ext4_xattr_ibody_header`` that is 4 bytes long:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - h\_magic
+     - Magic number for identification, 0xEA020000. This value is set by the
+       Linux driver, though e2fsprogs doesn't seem to check it(?)
+
+The beginning of an extended attribute block is in
+``struct ext4_xattr_header``, which is 32 bytes long:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - h\_magic
+     - Magic number for identification, 0xEA020000.
+   * - 0x4
+     - \_\_le32
+     - h\_refcount
+     - Reference count.
+   * - 0x8
+     - \_\_le32
+     - h\_blocks
+     - Number of disk blocks used.
+   * - 0xC
+     - \_\_le32
+     - h\_hash
+     - Hash value of all attributes.
+   * - 0x10
+     - \_\_le32
+     - h\_checksum
+     - Checksum of the extended attribute block.
+   * - 0x14
+     - \_\_u32
+     - h\_reserved[2]
+     - Zero.
+
+The checksum is calculated against the FS UUID, the 64-bit block number
+of the extended attribute block, and the entire block (header +
+entries).
+
+Following the ``struct ext4_xattr_header`` or
+``struct ext4_xattr_ibody_header`` is an array of
+``struct ext4_xattr_entry``; each of these entries is at least 16 bytes
+long. When stored in an external block, the ``struct ext4_xattr_entry``
+entries must be stored in sorted order. The sort order is
+``e_name_index``, then ``e_name_len``, and finally ``e_name``.
+Attributes stored inside an inode do not need be stored in sorted order.
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Description
+   * - 0x0
+     - \_\_u8
+     - e\_name\_len
+     - Length of name.
+   * - 0x1
+     - \_\_u8
+     - e\_name\_index
+     - Attribute name index. There is a discussion of this below.
+   * - 0x2
+     - \_\_le16
+     - e\_value\_offs
+     - Location of this attribute's value on the disk block where it is stored.
+       Multiple attributes can share the same value. For an inode attribute
+       this value is relative to the start of the first entry; for a block this
+       value is relative to the start of the block (i.e. the header).
+   * - 0x4
+     - \_\_le32
+     - e\_value\_inum
+     - The inode where the value is stored. Zero indicates the value is in the
+       same block as this entry. This field is only used if the
+       INCOMPAT\_EA\_INODE feature is enabled.
+   * - 0x8
+     - \_\_le32
+     - e\_value\_size
+     - Length of attribute value.
+   * - 0xC
+     - \_\_le32
+     - e\_hash
+     - Hash value of attribute name and attribute value. The kernel doesn't
+       update the hash for in-inode attributes, so for that case this value
+       must be zero, because e2fsck validates any non-zero hash regardless of
+       where the xattr lives.
+   * - 0x10
+     - char
+     - e\_name[e\_name\_len]
+     - Attribute name. Does not include trailing NULL.
+
+Attribute values can follow the end of the entry table. There appears to
+be a requirement that they be aligned to 4-byte boundaries. The values
+are stored starting at the end of the block and grow towards the
+xattr\_header/xattr\_entry table. When the two collide, the overflow is
+put into a separate disk block. If the disk block fills up, the
+filesystem returns -ENOSPC.
+
+The first four fields of the ``ext4_xattr_entry`` are set to zero to
+mark the end of the key list.
+
+Attribute Name Indices
+~~~~~~~~~~~~~~~~~~~~~~
+
+Logically speaking, extended attributes are a series of key=value pairs.
+The keys are assumed to be NULL-terminated strings. To reduce the amount
+of on-disk space that the keys consume, the beginning of the key string
+is matched against the attribute name index. If a match is found, the
+attribute name index field is set, and matching string is removed from
+the key name. Here is a map of name index values to key prefixes:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Name Index
+     - Key Prefix
+   * - 0
+     - (no prefix)
+   * - 1
+     - “user.”
+   * - 2
+     - “system.posix\_acl\_access”
+   * - 3
+     - “system.posix\_acl\_default”
+   * - 4
+     - “trusted.”
+   * - 6
+     - “security.”
+   * - 7
+     - “system.” (inline\_data only?)
+   * - 8
+     - “system.richacl” (SuSE kernels only?)
+
+For example, if the attribute key is “user.fubar”, the attribute name
+index is set to 1 and the “fubar” name is recorded on disk.
+
+POSIX ACLs
+~~~~~~~~~~
+
+POSIX ACLs are stored in a reduced version of the Linux kernel (and
+libacl's) internal ACL format. The key difference is that the version
+number is different (1) and the ``e_id`` field is only stored for named
+user and group ACLs.
diff --git a/Documentation/filesystems/ext4/ondisk/bigalloc.rst b/Documentation/filesystems/ext4/ondisk/bigalloc.rst
new file mode 100644
index 000000000..c6d885575
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/bigalloc.rst
@@ -0,0 +1,22 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Bigalloc
+--------
+
+At the moment, the default size of a block is 4KiB, which is a commonly
+supported page size on most MMU-capable hardware. This is fortunate, as
+ext4 code is not prepared to handle the case where the block size
+exceeds the page size. However, for a filesystem of mostly huge files,
+it is desirable to be able to allocate disk blocks in units of multiple
+blocks to reduce both fragmentation and metadata overhead. The
+`bigalloc <Bigalloc>`__ feature provides exactly this ability. The
+administrator can set a block cluster size at mkfs time (which is stored
+in the s\_log\_cluster\_size field in the superblock); from then on, the
+block bitmaps track clusters, not individual blocks. This means that
+block groups can be several gigabytes in size (instead of just 128MiB);
+however, the minimum allocation unit becomes a cluster, not a block,
+even for directories. TaoBao had a patchset to extend the “use units of
+clusters instead of blocks” to the extent tree, though it is not clear
+where those patches went-- they eventually morphed into “extent tree v2”
+but that code has not landed as of May 2015.
+
diff --git a/Documentation/filesystems/ext4/ondisk/bitmaps.rst b/Documentation/filesystems/ext4/ondisk/bitmaps.rst
new file mode 100644
index 000000000..c7546dbc1
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/bitmaps.rst
@@ -0,0 +1,28 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Block and inode Bitmaps
+-----------------------
+
+The data block bitmap tracks the usage of data blocks within the block
+group.
+
+The inode bitmap records which entries in the inode table are in use.
+
+As with most bitmaps, one bit represents the usage status of one data
+block or inode table entry. This implies a block group size of 8 \*
+number\_of\_bytes\_in\_a\_logical\_block.
+
+NOTE: If ``BLOCK_UNINIT`` is set for a given block group, various parts
+of the kernel and e2fsprogs code pretends that the block bitmap contains
+zeros (i.e. all blocks in the group are free). However, it is not
+necessarily the case that no blocks are in use -- if ``meta_bg`` is set,
+the bitmaps and group descriptor live inside the group. Unfortunately,
+ext2fs\_test\_block\_bitmap2() will return '0' for those locations,
+which produces confusing debugfs output.
+
+Inode Table
+-----------
+Inode tables are statically allocated at mkfs time.  Each block group
+descriptor points to the start of the table, and the superblock records
+the number of inodes per group.  See the section on inodes for more
+information.
diff --git a/Documentation/filesystems/ext4/ondisk/blockgroup.rst b/Documentation/filesystems/ext4/ondisk/blockgroup.rst
new file mode 100644
index 000000000..baf888e4c
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/blockgroup.rst
@@ -0,0 +1,135 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Layout
+------
+
+The layout of a standard block group is approximately as follows (each
+of these fields is discussed in a separate section below):
+
+.. list-table::
+   :widths: 1 1 1 1 1 1 1 1
+   :header-rows: 1
+
+   * - Group 0 Padding
+     - ext4 Super Block
+     - Group Descriptors
+     - Reserved GDT Blocks
+     - Data Block Bitmap
+     - inode Bitmap
+     - inode Table
+     - Data Blocks
+   * - 1024 bytes
+     - 1 block
+     - many blocks
+     - many blocks
+     - 1 block
+     - 1 block
+     - many blocks
+     - many more blocks
+
+For the special case of block group 0, the first 1024 bytes are unused,
+to allow for the installation of x86 boot sectors and other oddities.
+The superblock will start at offset 1024 bytes, whichever block that
+happens to be (usually 0). However, if for some reason the block size =
+1024, then block 0 is marked in use and the superblock goes in block 1.
+For all other block groups, there is no padding.
+
+The ext4 driver primarily works with the superblock and the group
+descriptors that are found in block group 0. Redundant copies of the
+superblock and group descriptors are written to some of the block groups
+across the disk in case the beginning of the disk gets trashed, though
+not all block groups necessarily host a redundant copy (see following
+paragraph for more details). If the group does not have a redundant
+copy, the block group begins with the data block bitmap. Note also that
+when the filesystem is freshly formatted, mkfs will allocate “reserve
+GDT block” space after the block group descriptors and before the start
+of the block bitmaps to allow for future expansion of the filesystem. By
+default, a filesystem is allowed to increase in size by a factor of
+1024x over the original filesystem size.
+
+The location of the inode table is given by ``grp.bg_inode_table_*``. It
+is continuous range of blocks large enough to contain
+``sb.s_inodes_per_group * sb.s_inode_size`` bytes.
+
+As for the ordering of items in a block group, it is generally
+established that the super block and the group descriptor table, if
+present, will be at the beginning of the block group. The bitmaps and
+the inode table can be anywhere, and it is quite possible for the
+bitmaps to come after the inode table, or for both to be in different
+groups (flex\_bg). Leftover space is used for file data blocks, indirect
+block maps, extent tree blocks, and extended attributes.
+
+Flexible Block Groups
+---------------------
+
+Starting in ext4, there is a new feature called flexible block groups
+(flex\_bg). In a flex\_bg, several block groups are tied together as one
+logical block group; the bitmap spaces and the inode table space in the
+first block group of the flex\_bg are expanded to include the bitmaps
+and inode tables of all other block groups in the flex\_bg. For example,
+if the flex\_bg size is 4, then group 0 will contain (in order) the
+superblock, group descriptors, data block bitmaps for groups 0-3, inode
+bitmaps for groups 0-3, inode tables for groups 0-3, and the remaining
+space in group 0 is for file data. The effect of this is to group the
+block metadata close together for faster loading, and to enable large
+files to be continuous on disk. Backup copies of the superblock and
+group descriptors are always at the beginning of block groups, even if
+flex\_bg is enabled. The number of block groups that make up a flex\_bg
+is given by 2 ^ ``sb.s_log_groups_per_flex``.
+
+Meta Block Groups
+-----------------
+
+Without the option META\_BG, for safety concerns, all block group
+descriptors copies are kept in the first block group. Given the default
+128MiB(2^27 bytes) block group size and 64-byte group descriptors, ext4
+can have at most 2^27/64 = 2^21 block groups. This limits the entire
+filesystem size to 2^21 ∗ 2^27 = 2^48bytes or 256TiB.
+
+The solution to this problem is to use the metablock group feature
+(META\_BG), which is already in ext3 for all 2.6 releases. With the
+META\_BG feature, ext4 filesystems are partitioned into many metablock
+groups. Each metablock group is a cluster of block groups whose group
+descriptor structures can be stored in a single disk block. For ext4
+filesystems with 4 KB block size, a single metablock group partition
+includes 64 block groups, or 8 GiB of disk space. The metablock group
+feature moves the location of the group descriptors from the congested
+first block group of the whole filesystem into the first group of each
+metablock group itself. The backups are in the second and last group of
+each metablock group. This increases the 2^21 maximum block groups limit
+to the hard limit 2^32, allowing support for a 512PiB filesystem.
+
+The change in the filesystem format replaces the current scheme where
+the superblock is followed by a variable-length set of block group
+descriptors. Instead, the superblock and a single block group descriptor
+block is placed at the beginning of the first, second, and last block
+groups in a meta-block group. A meta-block group is a collection of
+block groups which can be described by a single block group descriptor
+block. Since the size of the block group descriptor structure is 32
+bytes, a meta-block group contains 32 block groups for filesystems with
+a 1KB block size, and 128 block groups for filesystems with a 4KB
+blocksize. Filesystems can either be created using this new block group
+descriptor layout, or existing filesystems can be resized on-line, and
+the field s\_first\_meta\_bg in the superblock will indicate the first
+block group using this new layout.
+
+Please see an important note about ``BLOCK_UNINIT`` in the section about
+block and inode bitmaps.
+
+Lazy Block Group Initialization
+-------------------------------
+
+A new feature for ext4 are three block group descriptor flags that
+enable mkfs to skip initializing other parts of the block group
+metadata. Specifically, the INODE\_UNINIT and BLOCK\_UNINIT flags mean
+that the inode and block bitmaps for that group can be calculated and
+therefore the on-disk bitmap blocks are not initialized. This is
+generally the case for an empty block group or a block group containing
+only fixed-location block group metadata. The INODE\_ZEROED flag means
+that the inode table has been initialized; mkfs will unset this flag and
+rely on the kernel to initialize the inode tables in the background.
+
+By not writing zeroes to the bitmaps and inode table, mkfs time is
+reduced considerably. Note the feature flag is RO\_COMPAT\_GDT\_CSUM,
+but the dumpe2fs output prints this as “uninit\_bg”. They are the same
+thing.
diff --git a/Documentation/filesystems/ext4/ondisk/blockmap.rst b/Documentation/filesystems/ext4/ondisk/blockmap.rst
new file mode 100644
index 000000000..30e25750d
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/blockmap.rst
@@ -0,0 +1,49 @@
+.. SPDX-License-Identifier: GPL-2.0
+
++---------------------+------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------+
+| i.i\_block Offset   | Where It Points                                                                                                                                                                                                              |
++=====================+==============================================================================================================================================================================================================================+
+| 0 to 11             | Direct map to file blocks 0 to 11.                                                                                                                                                                                           |
++---------------------+------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------+
+| 12                  | Indirect block: (file blocks 12 to (``$block_size`` / 4) + 11, or 12 to 1035 if 4KiB blocks)                                                                                                                                 |
+|                     |                                                                                                                                                                                                                              |
+|                     | +------------------------------+--------------------------------------------------------------------+                                                                                                                        |
+|                     | | Indirect Block Offset        | Where It Points                                                    |                                                                                                                        |
+|                     | +==============================+====================================================================+                                                                                                                        |
+|                     | | 0 to (``$block_size`` / 4)   | Direct map to (``$block_size`` / 4) blocks (1024 if 4KiB blocks)   |                                                                                                                        |
+|                     | +------------------------------+--------------------------------------------------------------------+                                                                                                                        |
++---------------------+------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------+
+| 13                  | Double-indirect block: (file blocks ``$block_size``/4 + 12 to (``$block_size`` / 4) ^ 2 + (``$block_size`` / 4) + 11, or 1036 to 1049611 if 4KiB blocks)                                                                     |
+|                     |                                                                                                                                                                                                                              |
+|                     | +--------------------------------+---------------------------------------------------------------------------------------------------------+                                                                                 |
+|                     | | Double Indirect Block Offset   | Where It Points                                                                                         |                                                                                 |
+|                     | +================================+=========================================================================================================+                                                                                 |
+|                     | | 0 to (``$block_size`` / 4)     | Map to (``$block_size`` / 4) indirect blocks (1024 if 4KiB blocks)                                      |                                                                                 |
+|                     | |                                |                                                                                                         |                                                                                 |
+|                     | |                                | +------------------------------+--------------------------------------------------------------------+   |                                                                                 |
+|                     | |                                | | Indirect Block Offset        | Where It Points                                                    |   |                                                                                 |
+|                     | |                                | +==============================+====================================================================+   |                                                                                 |
+|                     | |                                | | 0 to (``$block_size`` / 4)   | Direct map to (``$block_size`` / 4) blocks (1024 if 4KiB blocks)   |   |                                                                                 |
+|                     | |                                | +------------------------------+--------------------------------------------------------------------+   |                                                                                 |
+|                     | +--------------------------------+---------------------------------------------------------------------------------------------------------+                                                                                 |
++---------------------+------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------+
+| 14                  | Triple-indirect block: (file blocks (``$block_size`` / 4) ^ 2 + (``$block_size`` / 4) + 12 to (``$block_size`` / 4) ^ 3 + (``$block_size`` / 4) ^ 2 + (``$block_size`` / 4) + 12, or 1049612 to 1074791436 if 4KiB blocks)   |
+|                     |                                                                                                                                                                                                                              |
+|                     | +--------------------------------+------------------------------------------------------------------------------------------------------------------------------------------------+                                          |
+|                     | | Triple Indirect Block Offset   | Where It Points                                                                                                                                |                                          |
+|                     | +================================+================================================================================================================================================+                                          |
+|                     | | 0 to (``$block_size`` / 4)     | Map to (``$block_size`` / 4) double indirect blocks (1024 if 4KiB blocks)                                                                      |                                          |
+|                     | |                                |                                                                                                                                                |                                          |
+|                     | |                                | +--------------------------------+---------------------------------------------------------------------------------------------------------+   |                                          |
+|                     | |                                | | Double Indirect Block Offset   | Where It Points                                                                                         |   |                                          |
+|                     | |                                | +================================+=========================================================================================================+   |                                          |
+|                     | |                                | | 0 to (``$block_size`` / 4)     | Map to (``$block_size`` / 4) indirect blocks (1024 if 4KiB blocks)                                      |   |                                          |
+|                     | |                                | |                                |                                                                                                         |   |                                          |
+|                     | |                                | |                                | +------------------------------+--------------------------------------------------------------------+   |   |                                          |
+|                     | |                                | |                                | | Indirect Block Offset        | Where It Points                                                    |   |   |                                          |
+|                     | |                                | |                                | +==============================+====================================================================+   |   |                                          |
+|                     | |                                | |                                | | 0 to (``$block_size`` / 4)   | Direct map to (``$block_size`` / 4) blocks (1024 if 4KiB blocks)   |   |   |                                          |
+|                     | |                                | |                                | +------------------------------+--------------------------------------------------------------------+   |   |                                          |
+|                     | |                                | +--------------------------------+---------------------------------------------------------------------------------------------------------+   |                                          |
+|                     | +--------------------------------+------------------------------------------------------------------------------------------------------------------------------------------------+                                          |
++---------------------+------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------------+
diff --git a/Documentation/filesystems/ext4/ondisk/blocks.rst b/Documentation/filesystems/ext4/ondisk/blocks.rst
new file mode 100644
index 000000000..73d4dc0f7
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/blocks.rst
@@ -0,0 +1,142 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Blocks
+------
+
+ext4 allocates storage space in units of “blocks”. A block is a group of
+sectors between 1KiB and 64KiB, and the number of sectors must be an
+integral power of 2. Blocks are in turn grouped into larger units called
+block groups. Block size is specified at mkfs time and typically is
+4KiB. You may experience mounting problems if block size is greater than
+page size (i.e. 64KiB blocks on a i386 which only has 4KiB memory
+pages). By default a filesystem can contain 2^32 blocks; if the '64bit'
+feature is enabled, then a filesystem can have 2^64 blocks.
+
+For 32-bit filesystems, limits are as follows:
+
+.. list-table::
+   :widths: 1 1 1 1 1
+   :header-rows: 1
+
+   * - Item
+     - 1KiB
+     - 2KiB
+     - 4KiB
+     - 64KiB
+   * - Blocks
+     - 2^32
+     - 2^32
+     - 2^32
+     - 2^32
+   * - Inodes
+     - 2^32
+     - 2^32
+     - 2^32
+     - 2^32
+   * - File System Size
+     - 4TiB
+     - 8TiB
+     - 16TiB
+     - 256PiB
+   * - Blocks Per Block Group
+     - 8,192
+     - 16,384
+     - 32,768
+     - 524,288
+   * - Inodes Per Block Group
+     - 8,192
+     - 16,384
+     - 32,768
+     - 524,288
+   * - Block Group Size
+     - 8MiB
+     - 32MiB
+     - 128MiB
+     - 32GiB
+   * - Blocks Per File, Extents
+     - 2^32
+     - 2^32
+     - 2^32
+     - 2^32
+   * - Blocks Per File, Block Maps
+     - 16,843,020
+     - 134,480,396
+     - 1,074,791,436
+     - 4,398,314,962,956 (really 2^32 due to field size limitations)
+   * - File Size, Extents
+     - 4TiB
+     - 8TiB
+     - 16TiB
+     - 256TiB
+   * - File Size, Block Maps
+     - 16GiB
+     - 256GiB
+     - 4TiB
+     - 256TiB
+
+For 64-bit filesystems, limits are as follows:
+
+.. list-table::
+   :widths: 1 1 1 1 1
+   :header-rows: 1
+
+   * - Item
+     - 1KiB
+     - 2KiB
+     - 4KiB
+     - 64KiB
+   * - Blocks
+     - 2^64
+     - 2^64
+     - 2^64
+     - 2^64
+   * - Inodes
+     - 2^32
+     - 2^32
+     - 2^32
+     - 2^32
+   * - File System Size
+     - 16ZiB
+     - 32ZiB
+     - 64ZiB
+     - 1YiB
+   * - Blocks Per Block Group
+     - 8,192
+     - 16,384
+     - 32,768
+     - 524,288
+   * - Inodes Per Block Group
+     - 8,192
+     - 16,384
+     - 32,768
+     - 524,288
+   * - Block Group Size
+     - 8MiB
+     - 32MiB
+     - 128MiB
+     - 32GiB
+   * - Blocks Per File, Extents
+     - 2^32
+     - 2^32
+     - 2^32
+     - 2^32
+   * - Blocks Per File, Block Maps
+     - 16,843,020
+     - 134,480,396
+     - 1,074,791,436
+     - 4,398,314,962,956 (really 2^32 due to field size limitations)
+   * - File Size, Extents
+     - 4TiB
+     - 8TiB
+     - 16TiB
+     - 256TiB
+   * - File Size, Block Maps
+     - 16GiB
+     - 256GiB
+     - 4TiB
+     - 256TiB
+
+Note: Files not using extents (i.e. files using block maps) must be
+placed within the first 2^32 blocks of a filesystem. Files with extents
+must be placed within the first 2^48 blocks of a filesystem. It's not
+clear what happens with larger filesystems.
diff --git a/Documentation/filesystems/ext4/ondisk/checksums.rst b/Documentation/filesystems/ext4/ondisk/checksums.rst
new file mode 100644
index 000000000..9d6a793b2
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/checksums.rst
@@ -0,0 +1,73 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Checksums
+---------
+
+Starting in early 2012, metadata checksums were added to all major ext4
+and jbd2 data structures. The associated feature flag is metadata\_csum.
+The desired checksum algorithm is indicated in the superblock, though as
+of October 2012 the only supported algorithm is crc32c. Some data
+structures did not have space to fit a full 32-bit checksum, so only the
+lower 16 bits are stored. Enabling the 64bit feature increases the data
+structure size so that full 32-bit checksums can be stored for many data
+structures. However, existing 32-bit filesystems cannot be extended to
+enable 64bit mode, at least not without the experimental resize2fs
+patches to do so.
+
+Existing filesystems can have checksumming added by running
+``tune2fs -O metadata_csum`` against the underlying device. If tune2fs
+encounters directory blocks that lack sufficient empty space to add a
+checksum, it will request that you run ``e2fsck -D`` to have the
+directories rebuilt with checksums. This has the added benefit of
+removing slack space from the directory files and rebalancing the htree
+indexes. If you \_ignore\_ this step, your directories will not be
+protected by a checksum!
+
+The following table describes the data elements that go into each type
+of checksum. The checksum function is whatever the superblock describes
+(crc32c as of October 2013) unless noted otherwise.
+
+.. list-table::
+   :widths: 1 1 4
+   :header-rows: 1
+
+   * - Metadata
+     - Length
+     - Ingredients
+   * - Superblock
+     - \_\_le32
+     - The entire superblock up to the checksum field. The UUID lives inside
+       the superblock.
+   * - MMP
+     - \_\_le32
+     - UUID + the entire MMP block up to the checksum field.
+   * - Extended Attributes
+     - \_\_le32
+     - UUID + the entire extended attribute block. The checksum field is set to
+       zero.
+   * - Directory Entries
+     - \_\_le32
+     - UUID + inode number + inode generation + the directory block up to the
+       fake entry enclosing the checksum field.
+   * - HTREE Nodes
+     - \_\_le32
+     - UUID + inode number + inode generation + all valid extents + HTREE tail.
+       The checksum field is set to zero.
+   * - Extents
+     - \_\_le32
+     - UUID + inode number + inode generation + the entire extent block up to
+       the checksum field.
+   * - Bitmaps
+     - \_\_le32 or \_\_le16
+     - UUID + the entire bitmap. Checksums are stored in the group descriptor,
+       and truncated if the group descriptor size is 32 bytes (i.e. ^64bit)
+   * - Inodes
+     - \_\_le32
+     - UUID + inode number + inode generation + the entire inode. The checksum
+       field is set to zero. Each inode has its own checksum.
+   * - Group Descriptors
+     - \_\_le16
+     - If metadata\_csum, then UUID + group number + the entire descriptor;
+       else if gdt\_csum, then crc16(UUID + group number + the entire
+       descriptor). In all cases, only the lower 16 bits are stored.
+
diff --git a/Documentation/filesystems/ext4/ondisk/directory.rst b/Documentation/filesystems/ext4/ondisk/directory.rst
new file mode 100644
index 000000000..8fcba68c2
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/directory.rst
@@ -0,0 +1,426 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Directory Entries
+-----------------
+
+In an ext4 filesystem, a directory is more or less a flat file that maps
+an arbitrary byte string (usually ASCII) to an inode number on the
+filesystem. There can be many directory entries across the filesystem
+that reference the same inode number--these are known as hard links, and
+that is why hard links cannot reference files on other filesystems. As
+such, directory entries are found by reading the data block(s)
+associated with a directory file for the particular directory entry that
+is desired.
+
+Linear (Classic) Directories
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+By default, each directory lists its entries in an “almost-linear”
+array. I write “almost” because it's not a linear array in the memory
+sense because directory entries are not split across filesystem blocks.
+Therefore, it is more accurate to say that a directory is a series of
+data blocks and that each block contains a linear array of directory
+entries. The end of each per-block array is signified by reaching the
+end of the block; the last entry in the block has a record length that
+takes it all the way to the end of the block. The end of the entire
+directory is of course signified by reaching the end of the file. Unused
+directory entries are signified by inode = 0. By default the filesystem
+uses ``struct ext4_dir_entry_2`` for directory entries unless the
+“filetype” feature flag is not set, in which case it uses
+``struct ext4_dir_entry``.
+
+The original directory entry format is ``struct ext4_dir_entry``, which
+is at most 263 bytes long, though on disk you'll need to reference
+``dirent.rec_len`` to know for sure.
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - inode
+     - Number of the inode that this directory entry points to.
+   * - 0x4
+     - \_\_le16
+     - rec\_len
+     - Length of this directory entry. Must be a multiple of 4.
+   * - 0x6
+     - \_\_le16
+     - name\_len
+     - Length of the file name.
+   * - 0x8
+     - char
+     - name[EXT4\_NAME\_LEN]
+     - File name.
+
+Since file names cannot be longer than 255 bytes, the new directory
+entry format shortens the rec\_len field and uses the space for a file
+type flag, probably to avoid having to load every inode during directory
+tree traversal. This format is ``ext4_dir_entry_2``, which is at most
+263 bytes long, though on disk you'll need to reference
+``dirent.rec_len`` to know for sure.
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - inode
+     - Number of the inode that this directory entry points to.
+   * - 0x4
+     - \_\_le16
+     - rec\_len
+     - Length of this directory entry.
+   * - 0x6
+     - \_\_u8
+     - name\_len
+     - Length of the file name.
+   * - 0x7
+     - \_\_u8
+     - file\_type
+     - File type code, see ftype_ table below.
+   * - 0x8
+     - char
+     - name[EXT4\_NAME\_LEN]
+     - File name.
+
+.. _ftype:
+
+The directory file type is one of the following values:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x0
+     - Unknown.
+   * - 0x1
+     - Regular file.
+   * - 0x2
+     - Directory.
+   * - 0x3
+     - Character device file.
+   * - 0x4
+     - Block device file.
+   * - 0x5
+     - FIFO.
+   * - 0x6
+     - Socket.
+   * - 0x7
+     - Symbolic link.
+
+In order to add checksums to these classic directory blocks, a phony
+``struct ext4_dir_entry`` is placed at the end of each leaf block to
+hold the checksum. The directory entry is 12 bytes long. The inode
+number and name\_len fields are set to zero to fool old software into
+ignoring an apparently empty directory entry, and the checksum is stored
+in the place where the name normally goes. The structure is
+``struct ext4_dir_entry_tail``:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - det\_reserved\_zero1
+     - Inode number, which must be zero.
+   * - 0x4
+     - \_\_le16
+     - det\_rec\_len
+     - Length of this directory entry, which must be 12.
+   * - 0x6
+     - \_\_u8
+     - det\_reserved\_zero2
+     - Length of the file name, which must be zero.
+   * - 0x7
+     - \_\_u8
+     - det\_reserved\_ft
+     - File type, which must be 0xDE.
+   * - 0x8
+     - \_\_le32
+     - det\_checksum
+     - Directory leaf block checksum.
+
+The leaf directory block checksum is calculated against the FS UUID, the
+directory's inode number, the directory's inode generation number, and
+the entire directory entry block up to (but not including) the fake
+directory entry.
+
+Hash Tree Directories
+~~~~~~~~~~~~~~~~~~~~~
+
+A linear array of directory entries isn't great for performance, so a
+new feature was added to ext3 to provide a faster (but peculiar)
+balanced tree keyed off a hash of the directory entry name. If the
+EXT4\_INDEX\_FL (0x1000) flag is set in the inode, this directory uses a
+hashed btree (htree) to organize and find directory entries. For
+backwards read-only compatibility with ext2, this tree is actually
+hidden inside the directory file, masquerading as “empty” directory data
+blocks! It was stated previously that the end of the linear directory
+entry table was signified with an entry pointing to inode 0; this is
+(ab)used to fool the old linear-scan algorithm into thinking that the
+rest of the directory block is empty so that it moves on.
+
+The root of the tree always lives in the first data block of the
+directory. By ext2 custom, the '.' and '..' entries must appear at the
+beginning of this first block, so they are put here as two
+``struct ext4_dir_entry_2``\ s and not stored in the tree. The rest of
+the root node contains metadata about the tree and finally a hash->block
+map to find nodes that are lower in the htree. If
+``dx_root.info.indirect_levels`` is non-zero then the htree has two
+levels; the data block pointed to by the root node's map is an interior
+node, which is indexed by a minor hash. Interior nodes in this tree
+contains a zeroed out ``struct ext4_dir_entry_2`` followed by a
+minor\_hash->block map to find leafe nodes. Leaf nodes contain a linear
+array of all ``struct ext4_dir_entry_2``; all of these entries
+(presumably) hash to the same value. If there is an overflow, the
+entries simply overflow into the next leaf node, and the
+least-significant bit of the hash (in the interior node map) that gets
+us to this next leaf node is set.
+
+To traverse the directory as a htree, the code calculates the hash of
+the desired file name and uses it to find the corresponding block
+number. If the tree is flat, the block is a linear array of directory
+entries that can be searched; otherwise, the minor hash of the file name
+is computed and used against this second block to find the corresponding
+third block number. That third block number will be a linear array of
+directory entries.
+
+To traverse the directory as a linear array (such as the old code does),
+the code simply reads every data block in the directory. The blocks used
+for the htree will appear to have no entries (aside from '.' and '..')
+and so only the leaf nodes will appear to have any interesting content.
+
+The root of the htree is in ``struct dx_root``, which is the full length
+of a data block:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - dot.inode
+     - inode number of this directory.
+   * - 0x4
+     - \_\_le16
+     - dot.rec\_len
+     - Length of this record, 12.
+   * - 0x6
+     - u8
+     - dot.name\_len
+     - Length of the name, 1.
+   * - 0x7
+     - u8
+     - dot.file\_type
+     - File type of this entry, 0x2 (directory) (if the feature flag is set).
+   * - 0x8
+     - char
+     - dot.name[4]
+     - “.\\0\\0\\0”
+   * - 0xC
+     - \_\_le32
+     - dotdot.inode
+     - inode number of parent directory.
+   * - 0x10
+     - \_\_le16
+     - dotdot.rec\_len
+     - block\_size - 12. The record length is long enough to cover all htree
+       data.
+   * - 0x12
+     - u8
+     - dotdot.name\_len
+     - Length of the name, 2.
+   * - 0x13
+     - u8
+     - dotdot.file\_type
+     - File type of this entry, 0x2 (directory) (if the feature flag is set).
+   * - 0x14
+     - char
+     - dotdot\_name[4]
+     - “..\\0\\0”
+   * - 0x18
+     - \_\_le32
+     - struct dx\_root\_info.reserved\_zero
+     - Zero.
+   * - 0x1C
+     - u8
+     - struct dx\_root\_info.hash\_version
+     - Hash type, see dirhash_ table below.
+   * - 0x1D
+     - u8
+     - struct dx\_root\_info.info\_length
+     - Length of the tree information, 0x8.
+   * - 0x1E
+     - u8
+     - struct dx\_root\_info.indirect\_levels
+     - Depth of the htree. Cannot be larger than 3 if the INCOMPAT\_LARGEDIR
+       feature is set; cannot be larger than 2 otherwise.
+   * - 0x1F
+     - u8
+     - struct dx\_root\_info.unused\_flags
+     -
+   * - 0x20
+     - \_\_le16
+     - limit
+     - Maximum number of dx\_entries that can follow this header, plus 1 for
+       the header itself.
+   * - 0x22
+     - \_\_le16
+     - count
+     - Actual number of dx\_entries that follow this header, plus 1 for the
+       header itself.
+   * - 0x24
+     - \_\_le32
+     - block
+     - The block number (within the directory file) that goes with hash=0.
+   * - 0x28
+     - struct dx\_entry
+     - entries[0]
+     - As many 8-byte ``struct dx_entry`` as fits in the rest of the data block.
+
+.. _dirhash:
+
+The directory hash is one of the following values:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x0
+     - Legacy.
+   * - 0x1
+     - Half MD4.
+   * - 0x2
+     - Tea.
+   * - 0x3
+     - Legacy, unsigned.
+   * - 0x4
+     - Half MD4, unsigned.
+   * - 0x5
+     - Tea, unsigned.
+
+Interior nodes of an htree are recorded as ``struct dx_node``, which is
+also the full length of a data block:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - fake.inode
+     - Zero, to make it look like this entry is not in use.
+   * - 0x4
+     - \_\_le16
+     - fake.rec\_len
+     - The size of the block, in order to hide all of the dx\_node data.
+   * - 0x6
+     - u8
+     - name\_len
+     - Zero. There is no name for this “unused” directory entry.
+   * - 0x7
+     - u8
+     - file\_type
+     - Zero. There is no file type for this “unused” directory entry.
+   * - 0x8
+     - \_\_le16
+     - limit
+     - Maximum number of dx\_entries that can follow this header, plus 1 for
+       the header itself.
+   * - 0xA
+     - \_\_le16
+     - count
+     - Actual number of dx\_entries that follow this header, plus 1 for the
+       header itself.
+   * - 0xE
+     - \_\_le32
+     - block
+     - The block number (within the directory file) that goes with the lowest
+       hash value of this block. This value is stored in the parent block.
+   * - 0x12
+     - struct dx\_entry
+     - entries[0]
+     - As many 8-byte ``struct dx_entry`` as fits in the rest of the data block.
+
+The hash maps that exist in both ``struct dx_root`` and
+``struct dx_node`` are recorded as ``struct dx_entry``, which is 8 bytes
+long:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - hash
+     - Hash code.
+   * - 0x4
+     - \_\_le32
+     - block
+     - Block number (within the directory file, not filesystem blocks) of the
+       next node in the htree.
+
+(If you think this is all quite clever and peculiar, so does the
+author.)
+
+If metadata checksums are enabled, the last 8 bytes of the directory
+block (precisely the length of one dx\_entry) are used to store a
+``struct dx_tail``, which contains the checksum. The ``limit`` and
+``count`` entries in the dx\_root/dx\_node structures are adjusted as
+necessary to fit the dx\_tail into the block. If there is no space for
+the dx\_tail, the user is notified to run e2fsck -D to rebuild the
+directory index (which will ensure that there's space for the checksum.
+The dx\_tail structure is 8 bytes long and looks like this:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Description
+   * - 0x0
+     - u32
+     - dt\_reserved
+     - Zero.
+   * - 0x4
+     - \_\_le32
+     - dt\_checksum
+     - Checksum of the htree directory block.
+
+The checksum is calculated against the FS UUID, the htree index header
+(dx\_root or dx\_node), all of the htree indices (dx\_entry) that are in
+use, and the tail block (dx\_tail).
diff --git a/Documentation/filesystems/ext4/ondisk/dynamic.rst b/Documentation/filesystems/ext4/ondisk/dynamic.rst
new file mode 100644
index 000000000..bb0c84333
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/dynamic.rst
@@ -0,0 +1,12 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Dynamic Structures
+==================
+
+Dynamic metadata are created on the fly when files and blocks are
+allocated to files.
+
+.. include:: inodes.rst
+.. include:: ifork.rst
+.. include:: directory.rst
+.. include:: attributes.rst
diff --git a/Documentation/filesystems/ext4/ondisk/eainode.rst b/Documentation/filesystems/ext4/ondisk/eainode.rst
new file mode 100644
index 000000000..ecc0d01a0
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/eainode.rst
@@ -0,0 +1,18 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Large Extended Attribute Values
+-------------------------------
+
+To enable ext4 to store extended attribute values that do not fit in the
+inode or in the single extended attribute block attached to an inode,
+the EA\_INODE feature allows us to store the value in the data blocks of
+a regular file inode. This “EA inode” is linked only from the extended
+attribute name index and must not appear in a directory entry. The
+inode's i\_atime field is used to store a checksum of the xattr value;
+and i\_ctime/i\_version store a 64-bit reference count, which enables
+sharing of large xattr values between multiple owning inodes. For
+backward compatibility with older versions of this feature, the
+i\_mtime/i\_generation *may* store a back-reference to the inode number
+and i\_generation of the **one** owning inode (in cases where the EA
+inode is not referenced by multiple inodes) to verify that the EA inode
+is the correct one being accessed.
diff --git a/Documentation/filesystems/ext4/ondisk/globals.rst b/Documentation/filesystems/ext4/ondisk/globals.rst
new file mode 100644
index 000000000..368bf7662
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/globals.rst
@@ -0,0 +1,13 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Global Structures
+=================
+
+The filesystem is sharded into a number of block groups, each of which
+have static metadata at fixed locations.
+
+.. include:: super.rst
+.. include:: group_descr.rst
+.. include:: bitmaps.rst
+.. include:: mmp.rst
+.. include:: journal.rst
diff --git a/Documentation/filesystems/ext4/ondisk/group_descr.rst b/Documentation/filesystems/ext4/ondisk/group_descr.rst
new file mode 100644
index 000000000..759827e5d
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/group_descr.rst
@@ -0,0 +1,170 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Block Group Descriptors
+-----------------------
+
+Each block group on the filesystem has one of these descriptors
+associated with it. As noted in the Layout section above, the group
+descriptors (if present) are the second item in the block group. The
+standard configuration is for each block group to contain a full copy of
+the block group descriptor table unless the sparse\_super feature flag
+is set.
+
+Notice how the group descriptor records the location of both bitmaps and
+the inode table (i.e. they can float). This means that within a block
+group, the only data structures with fixed locations are the superblock
+and the group descriptor table. The flex\_bg mechanism uses this
+property to group several block groups into a flex group and lay out all
+of the groups' bitmaps and inode tables into one long run in the first
+group of the flex group.
+
+If the meta\_bg feature flag is set, then several block groups are
+grouped together into a meta group. Note that in the meta\_bg case,
+however, the first and last two block groups within the larger meta
+group contain only group descriptors for the groups inside the meta
+group.
+
+flex\_bg and meta\_bg do not appear to be mutually exclusive features.
+
+In ext2, ext3, and ext4 (when the 64bit feature is not enabled), the
+block group descriptor was only 32 bytes long and therefore ends at
+bg\_checksum. On an ext4 filesystem with the 64bit feature enabled, the
+block group descriptor expands to at least the 64 bytes described below;
+the size is stored in the superblock.
+
+If gdt\_csum is set and metadata\_csum is not set, the block group
+checksum is the crc16 of the FS UUID, the group number, and the group
+descriptor structure. If metadata\_csum is set, then the block group
+checksum is the lower 16 bits of the checksum of the FS UUID, the group
+number, and the group descriptor structure. Both block and inode bitmap
+checksums are calculated against the FS UUID, the group number, and the
+entire bitmap.
+
+The block group descriptor is laid out in ``struct ext4_group_desc``.
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - bg\_block\_bitmap\_lo
+     - Lower 32-bits of location of block bitmap.
+   * - 0x4
+     - \_\_le32
+     - bg\_inode\_bitmap\_lo
+     - Lower 32-bits of location of inode bitmap.
+   * - 0x8
+     - \_\_le32
+     - bg\_inode\_table\_lo
+     - Lower 32-bits of location of inode table.
+   * - 0xC
+     - \_\_le16
+     - bg\_free\_blocks\_count\_lo
+     - Lower 16-bits of free block count.
+   * - 0xE
+     - \_\_le16
+     - bg\_free\_inodes\_count\_lo
+     - Lower 16-bits of free inode count.
+   * - 0x10
+     - \_\_le16
+     - bg\_used\_dirs\_count\_lo
+     - Lower 16-bits of directory count.
+   * - 0x12
+     - \_\_le16
+     - bg\_flags
+     - Block group flags. See the bgflags_ table below.
+   * - 0x14
+     - \_\_le32
+     - bg\_exclude\_bitmap\_lo
+     - Lower 32-bits of location of snapshot exclusion bitmap.
+   * - 0x18
+     - \_\_le16
+     - bg\_block\_bitmap\_csum\_lo
+     - Lower 16-bits of the block bitmap checksum.
+   * - 0x1A
+     - \_\_le16
+     - bg\_inode\_bitmap\_csum\_lo
+     - Lower 16-bits of the inode bitmap checksum.
+   * - 0x1C
+     - \_\_le16
+     - bg\_itable\_unused\_lo
+     - Lower 16-bits of unused inode count. If set, we needn't scan past the
+       ``(sb.s_inodes_per_group - gdt.bg_itable_unused)``\ th entry in the
+       inode table for this group.
+   * - 0x1E
+     - \_\_le16
+     - bg\_checksum
+     - Group descriptor checksum; crc16(sb\_uuid+group+desc) if the
+       RO\_COMPAT\_GDT\_CSUM feature is set, or crc32c(sb\_uuid+group\_desc) &
+       0xFFFF if the RO\_COMPAT\_METADATA\_CSUM feature is set.
+   * -
+     -
+     -
+     - These fields only exist if the 64bit feature is enabled and s_desc_size
+       > 32.
+   * - 0x20
+     - \_\_le32
+     - bg\_block\_bitmap\_hi
+     - Upper 32-bits of location of block bitmap.
+   * - 0x24
+     - \_\_le32
+     - bg\_inode\_bitmap\_hi
+     - Upper 32-bits of location of inodes bitmap.
+   * - 0x28
+     - \_\_le32
+     - bg\_inode\_table\_hi
+     - Upper 32-bits of location of inodes table.
+   * - 0x2C
+     - \_\_le16
+     - bg\_free\_blocks\_count\_hi
+     - Upper 16-bits of free block count.
+   * - 0x2E
+     - \_\_le16
+     - bg\_free\_inodes\_count\_hi
+     - Upper 16-bits of free inode count.
+   * - 0x30
+     - \_\_le16
+     - bg\_used\_dirs\_count\_hi
+     - Upper 16-bits of directory count.
+   * - 0x32
+     - \_\_le16
+     - bg\_itable\_unused\_hi
+     - Upper 16-bits of unused inode count.
+   * - 0x34
+     - \_\_le32
+     - bg\_exclude\_bitmap\_hi
+     - Upper 32-bits of location of snapshot exclusion bitmap.
+   * - 0x38
+     - \_\_le16
+     - bg\_block\_bitmap\_csum\_hi
+     - Upper 16-bits of the block bitmap checksum.
+   * - 0x3A
+     - \_\_le16
+     - bg\_inode\_bitmap\_csum\_hi
+     - Upper 16-bits of the inode bitmap checksum.
+   * - 0x3C
+     - \_\_u32
+     - bg\_reserved
+     - Padding to 64 bytes.
+
+.. _bgflags:
+
+Block group flags can be any combination of the following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x1
+     - inode table and bitmap are not initialized (EXT4\_BG\_INODE\_UNINIT).
+   * - 0x2
+     - block bitmap is not initialized (EXT4\_BG\_BLOCK\_UNINIT).
+   * - 0x4
+     - inode table is zeroed (EXT4\_BG\_INODE\_ZEROED).
diff --git a/Documentation/filesystems/ext4/ondisk/ifork.rst b/Documentation/filesystems/ext4/ondisk/ifork.rst
new file mode 100644
index 000000000..5dbe3b2b1
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/ifork.rst
@@ -0,0 +1,194 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+The Contents of inode.i\_block
+------------------------------
+
+Depending on the type of file an inode describes, the 60 bytes of
+storage in ``inode.i_block`` can be used in different ways. In general,
+regular files and directories will use it for file block indexing
+information, and special files will use it for special purposes.
+
+Symbolic Links
+~~~~~~~~~~~~~~
+
+The target of a symbolic link will be stored in this field if the target
+string is less than 60 bytes long. Otherwise, either extents or block
+maps will be used to allocate data blocks to store the link target.
+
+Direct/Indirect Block Addressing
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+In ext2/3, file block numbers were mapped to logical block numbers by
+means of an (up to) three level 1-1 block map. To find the logical block
+that stores a particular file block, the code would navigate through
+this increasingly complicated structure. Notice that there is neither a
+magic number nor a checksum to provide any level of confidence that the
+block isn't full of garbage.
+
+.. ifconfig:: builder != 'latex'
+
+   .. include:: blockmap.rst
+
+.. ifconfig:: builder == 'latex'
+
+   [Table omitted because LaTeX doesn't support nested tables.]
+
+Note that with this block mapping scheme, it is necessary to fill out a
+lot of mapping data even for a large contiguous file! This inefficiency
+led to the creation of the extent mapping scheme, discussed below.
+
+Notice also that a file using this mapping scheme cannot be placed
+higher than 2^32 blocks.
+
+Extent Tree
+~~~~~~~~~~~
+
+In ext4, the file to logical block map has been replaced with an extent
+tree. Under the old scheme, allocating a contiguous run of 1,000 blocks
+requires an indirect block to map all 1,000 entries; with extents, the
+mapping is reduced to a single ``struct ext4_extent`` with
+``ee_len = 1000``. If flex\_bg is enabled, it is possible to allocate
+very large files with a single extent, at a considerable reduction in
+metadata block use, and some improvement in disk efficiency. The inode
+must have the extents flag (0x80000) flag set for this feature to be in
+use.
+
+Extents are arranged as a tree. Each node of the tree begins with a
+``struct ext4_extent_header``. If the node is an interior node
+(``eh.eh_depth`` > 0), the header is followed by ``eh.eh_entries``
+instances of ``struct ext4_extent_idx``; each of these index entries
+points to a block containing more nodes in the extent tree. If the node
+is a leaf node (``eh.eh_depth == 0``), then the header is followed by
+``eh.eh_entries`` instances of ``struct ext4_extent``; these instances
+point to the file's data blocks. The root node of the extent tree is
+stored in ``inode.i_block``, which allows for the first four extents to
+be recorded without the use of extra metadata blocks.
+
+The extent tree header is recorded in ``struct ext4_extent_header``,
+which is 12 bytes long:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le16
+     - eh\_magic
+     - Magic number, 0xF30A.
+   * - 0x2
+     - \_\_le16
+     - eh\_entries
+     - Number of valid entries following the header.
+   * - 0x4
+     - \_\_le16
+     - eh\_max
+     - Maximum number of entries that could follow the header.
+   * - 0x6
+     - \_\_le16
+     - eh\_depth
+     - Depth of this extent node in the extent tree. 0 = this extent node
+       points to data blocks; otherwise, this extent node points to other
+       extent nodes. The extent tree can be at most 5 levels deep: a logical
+       block number can be at most ``2^32``, and the smallest ``n`` that
+       satisfies ``4*(((blocksize - 12)/12)^n) >= 2^32`` is 5.
+   * - 0x8
+     - \_\_le32
+     - eh\_generation
+     - Generation of the tree. (Used by Lustre, but not standard ext4).
+
+Internal nodes of the extent tree, also known as index nodes, are
+recorded as ``struct ext4_extent_idx``, and are 12 bytes long:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - ei\_block
+     - This index node covers file blocks from 'block' onward.
+   * - 0x4
+     - \_\_le32
+     - ei\_leaf\_lo
+     - Lower 32-bits of the block number of the extent node that is the next
+       level lower in the tree. The tree node pointed to can be either another
+       internal node or a leaf node, described below.
+   * - 0x8
+     - \_\_le16
+     - ei\_leaf\_hi
+     - Upper 16-bits of the previous field.
+   * - 0xA
+     - \_\_u16
+     - ei\_unused
+     -
+
+Leaf nodes of the extent tree are recorded as ``struct ext4_extent``,
+and are also 12 bytes long:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - ee\_block
+     - First file block number that this extent covers.
+   * - 0x4
+     - \_\_le16
+     - ee\_len
+     - Number of blocks covered by extent. If the value of this field is <=
+       32768, the extent is initialized. If the value of the field is > 32768,
+       the extent is uninitialized and the actual extent length is ``ee_len`` -
+       32768. Therefore, the maximum length of a initialized extent is 32768
+       blocks, and the maximum length of an uninitialized extent is 32767.
+   * - 0x6
+     - \_\_le16
+     - ee\_start\_hi
+     - Upper 16-bits of the block number to which this extent points.
+   * - 0x8
+     - \_\_le32
+     - ee\_start\_lo
+     - Lower 32-bits of the block number to which this extent points.
+
+Prior to the introduction of metadata checksums, the extent header +
+extent entries always left at least 4 bytes of unallocated space at the
+end of each extent tree data block (because (2^x % 12) >= 4). Therefore,
+the 32-bit checksum is inserted into this space. The 4 extents in the
+inode do not need checksumming, since the inode is already checksummed.
+The checksum is calculated against the FS UUID, the inode number, the
+inode generation, and the entire extent block leading up to (but not
+including) the checksum itself.
+
+``struct ext4_extent_tail`` is 4 bytes long:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - eb\_checksum
+     - Checksum of the extent block, crc32c(uuid+inum+igeneration+extentblock)
+
+Inline Data
+~~~~~~~~~~~
+
+If the inline data feature is enabled for the filesystem and the flag is
+set for the inode, it is possible that the first 60 bytes of the file
+data are stored here.
diff --git a/Documentation/filesystems/ext4/ondisk/index.rst b/Documentation/filesystems/ext4/ondisk/index.rst
new file mode 100644
index 000000000..f7d082c3a
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/index.rst
@@ -0,0 +1,9 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+==============================
+Data Structures and Algorithms
+==============================
+.. include:: about.rst
+.. include:: overview.rst
+.. include:: globals.rst
+.. include:: dynamic.rst
diff --git a/Documentation/filesystems/ext4/ondisk/inlinedata.rst b/Documentation/filesystems/ext4/ondisk/inlinedata.rst
new file mode 100644
index 000000000..d1075178c
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/inlinedata.rst
@@ -0,0 +1,37 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Inline Data
+-----------
+
+The inline data feature was designed to handle the case that a file's
+data is so tiny that it readily fits inside the inode, which
+(theoretically) reduces disk block consumption and reduces seeks. If the
+file is smaller than 60 bytes, then the data are stored inline in
+``inode.i_block``. If the rest of the file would fit inside the extended
+attribute space, then it might be found as an extended attribute
+“system.data” within the inode body (“ibody EA”). This of course
+constrains the amount of extended attributes one can attach to an inode.
+If the data size increases beyond i\_block + ibody EA, a regular block
+is allocated and the contents moved to that block.
+
+Pending a change to compact the extended attribute key used to store
+inline data, one ought to be able to store 160 bytes of data in a
+256-byte inode (as of June 2015, when i\_extra\_isize is 28). Prior to
+that, the limit was 156 bytes due to inefficient use of inode space.
+
+The inline data feature requires the presence of an extended attribute
+for “system.data”, even if the attribute value is zero length.
+
+Inline Directories
+~~~~~~~~~~~~~~~~~~
+
+The first four bytes of i\_block are the inode number of the parent
+directory. Following that is a 56-byte space for an array of directory
+entries; see ``struct ext4_dir_entry``. If there is a “system.data”
+attribute in the inode body, the EA value is an array of
+``struct ext4_dir_entry`` as well. Note that for inline directories, the
+i\_block and EA space are treated as separate dirent blocks; directory
+entries cannot span the two.
+
+Inline directory entries are not checksummed, as the inode checksum
+should protect all inline data contents.
diff --git a/Documentation/filesystems/ext4/ondisk/inodes.rst b/Documentation/filesystems/ext4/ondisk/inodes.rst
new file mode 100644
index 000000000..655ce898f
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/inodes.rst
@@ -0,0 +1,575 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Index Nodes
+-----------
+
+In a regular UNIX filesystem, the inode stores all the metadata
+pertaining to the file (time stamps, block maps, extended attributes,
+etc), not the directory entry. To find the information associated with a
+file, one must traverse the directory files to find the directory entry
+associated with a file, then load the inode to find the metadata for
+that file. ext4 appears to cheat (for performance reasons) a little bit
+by storing a copy of the file type (normally stored in the inode) in the
+directory entry. (Compare all this to FAT, which stores all the file
+information directly in the directory entry, but does not support hard
+links and is in general more seek-happy than ext4 due to its simpler
+block allocator and extensive use of linked lists.)
+
+The inode table is a linear array of ``struct ext4_inode``. The table is
+sized to have enough blocks to store at least
+``sb.s_inode_size * sb.s_inodes_per_group`` bytes. The number of the
+block group containing an inode can be calculated as
+``(inode_number - 1) / sb.s_inodes_per_group``, and the offset into the
+group's table is ``(inode_number - 1) % sb.s_inodes_per_group``. There
+is no inode 0.
+
+The inode checksum is calculated against the FS UUID, the inode number,
+and the inode structure itself.
+
+The inode table entry is laid out in ``struct ext4_inode``.
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le16
+     - i\_mode
+     - File mode. See the table i_mode_ below.
+   * - 0x2
+     - \_\_le16
+     - i\_uid
+     - Lower 16-bits of Owner UID.
+   * - 0x4
+     - \_\_le32
+     - i\_size\_lo
+     - Lower 32-bits of size in bytes.
+   * - 0x8
+     - \_\_le32
+     - i\_atime
+     - Last access time, in seconds since the epoch. However, if the EA\_INODE
+       inode flag is set, this inode stores an extended attribute value and
+       this field contains the checksum of the value.
+   * - 0xC
+     - \_\_le32
+     - i\_ctime
+     - Last inode change time, in seconds since the epoch. However, if the
+       EA\_INODE inode flag is set, this inode stores an extended attribute
+       value and this field contains the lower 32 bits of the attribute value's
+       reference count.
+   * - 0x10
+     - \_\_le32
+     - i\_mtime
+     - Last data modification time, in seconds since the epoch. However, if the
+       EA\_INODE inode flag is set, this inode stores an extended attribute
+       value and this field contains the number of the inode that owns the
+       extended attribute.
+   * - 0x14
+     - \_\_le32
+     - i\_dtime
+     - Deletion Time, in seconds since the epoch.
+   * - 0x18
+     - \_\_le16
+     - i\_gid
+     - Lower 16-bits of GID.
+   * - 0x1A
+     - \_\_le16
+     - i\_links\_count
+     - Hard link count. Normally, ext4 does not permit an inode to have more
+       than 65,000 hard links. This applies to files as well as directories,
+       which means that there cannot be more than 64,998 subdirectories in a
+       directory (each subdirectory's '..' entry counts as a hard link, as does
+       the '.' entry in the directory itself). With the DIR\_NLINK feature
+       enabled, ext4 supports more than 64,998 subdirectories by setting this
+       field to 1 to indicate that the number of hard links is not known.
+   * - 0x1C
+     - \_\_le32
+     - i\_blocks\_lo
+     - Lower 32-bits of “block” count. If the huge\_file feature flag is not
+       set on the filesystem, the file consumes ``i_blocks_lo`` 512-byte blocks
+       on disk. If huge\_file is set and EXT4\_HUGE\_FILE\_FL is NOT set in
+       ``inode.i_flags``, then the file consumes ``i_blocks_lo + (i_blocks_hi
+       << 32)`` 512-byte blocks on disk. If huge\_file is set and
+       EXT4\_HUGE\_FILE\_FL IS set in ``inode.i_flags``, then this file
+       consumes (``i_blocks_lo + i_blocks_hi`` << 32) filesystem blocks on
+       disk.
+   * - 0x20
+     - \_\_le32
+     - i\_flags
+     - Inode flags. See the table i_flags_ below.
+   * - 0x24
+     - 4 bytes
+     - i\_osd1
+     - See the table i_osd1_ for more details.
+   * - 0x28
+     - 60 bytes
+     - i\_block[EXT4\_N\_BLOCKS=15]
+     - Block map or extent tree. See the section “The Contents of inode.i\_block”.
+   * - 0x64
+     - \_\_le32
+     - i\_generation
+     - File version (for NFS).
+   * - 0x68
+     - \_\_le32
+     - i\_file\_acl\_lo
+     - Lower 32-bits of extended attribute block. ACLs are of course one of
+       many possible extended attributes; I think the name of this field is a
+       result of the first use of extended attributes being for ACLs.
+   * - 0x6C
+     - \_\_le32
+     - i\_size\_high / i\_dir\_acl
+     - Upper 32-bits of file/directory size. In ext2/3 this field was named
+       i\_dir\_acl, though it was usually set to zero and never used.
+   * - 0x70
+     - \_\_le32
+     - i\_obso\_faddr
+     - (Obsolete) fragment address.
+   * - 0x74
+     - 12 bytes
+     - i\_osd2
+     - See the table i_osd2_ for more details.
+   * - 0x80
+     - \_\_le16
+     - i\_extra\_isize
+     - Size of this inode - 128. Alternately, the size of the extended inode
+       fields beyond the original ext2 inode, including this field.
+   * - 0x82
+     - \_\_le16
+     - i\_checksum\_hi
+     - Upper 16-bits of the inode checksum.
+   * - 0x84
+     - \_\_le32
+     - i\_ctime\_extra
+     - Extra change time bits. This provides sub-second precision. See Inode
+       Timestamps section.
+   * - 0x88
+     - \_\_le32
+     - i\_mtime\_extra
+     - Extra modification time bits. This provides sub-second precision.
+   * - 0x8C
+     - \_\_le32
+     - i\_atime\_extra
+     - Extra access time bits. This provides sub-second precision.
+   * - 0x90
+     - \_\_le32
+     - i\_crtime
+     - File creation time, in seconds since the epoch.
+   * - 0x94
+     - \_\_le32
+     - i\_crtime\_extra
+     - Extra file creation time bits. This provides sub-second precision.
+   * - 0x98
+     - \_\_le32
+     - i\_version\_hi
+     - Upper 32-bits for version number.
+   * - 0x9C
+     - \_\_le32
+     - i\_projid
+     - Project ID.
+
+.. _i_mode:
+
+The ``i_mode`` value is a combination of the following flags:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x1
+     - S\_IXOTH (Others may execute)
+   * - 0x2
+     - S\_IWOTH (Others may write)
+   * - 0x4
+     - S\_IROTH (Others may read)
+   * - 0x8
+     - S\_IXGRP (Group members may execute)
+   * - 0x10
+     - S\_IWGRP (Group members may write)
+   * - 0x20
+     - S\_IRGRP (Group members may read)
+   * - 0x40
+     - S\_IXUSR (Owner may execute)
+   * - 0x80
+     - S\_IWUSR (Owner may write)
+   * - 0x100
+     - S\_IRUSR (Owner may read)
+   * - 0x200
+     - S\_ISVTX (Sticky bit)
+   * - 0x400
+     - S\_ISGID (Set GID)
+   * - 0x800
+     - S\_ISUID (Set UID)
+   * -
+     - These are mutually-exclusive file types:
+   * - 0x1000
+     - S\_IFIFO (FIFO)
+   * - 0x2000
+     - S\_IFCHR (Character device)
+   * - 0x4000
+     - S\_IFDIR (Directory)
+   * - 0x6000
+     - S\_IFBLK (Block device)
+   * - 0x8000
+     - S\_IFREG (Regular file)
+   * - 0xA000
+     - S\_IFLNK (Symbolic link)
+   * - 0xC000
+     - S\_IFSOCK (Socket)
+
+.. _i_flags:
+
+The ``i_flags`` field is a combination of these values:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x1
+     - This file requires secure deletion (EXT4\_SECRM\_FL). (not implemented)
+   * - 0x2
+     - This file should be preserved, should undeletion be desired
+       (EXT4\_UNRM\_FL). (not implemented)
+   * - 0x4
+     - File is compressed (EXT4\_COMPR\_FL). (not really implemented)
+   * - 0x8
+     - All writes to the file must be synchronous (EXT4\_SYNC\_FL).
+   * - 0x10
+     - File is immutable (EXT4\_IMMUTABLE\_FL).
+   * - 0x20
+     - File can only be appended (EXT4\_APPEND\_FL).
+   * - 0x40
+     - The dump(1) utility should not dump this file (EXT4\_NODUMP\_FL).
+   * - 0x80
+     - Do not update access time (EXT4\_NOATIME\_FL).
+   * - 0x100
+     - Dirty compressed file (EXT4\_DIRTY\_FL). (not used)
+   * - 0x200
+     - File has one or more compressed clusters (EXT4\_COMPRBLK\_FL). (not used)
+   * - 0x400
+     - Do not compress file (EXT4\_NOCOMPR\_FL). (not used)
+   * - 0x800
+     - Encrypted inode (EXT4\_ENCRYPT\_FL). This bit value previously was
+       EXT4\_ECOMPR\_FL (compression error), which was never used.
+   * - 0x1000
+     - Directory has hashed indexes (EXT4\_INDEX\_FL).
+   * - 0x2000
+     - AFS magic directory (EXT4\_IMAGIC\_FL).
+   * - 0x4000
+     - File data must always be written through the journal
+       (EXT4\_JOURNAL\_DATA\_FL).
+   * - 0x8000
+     - File tail should not be merged (EXT4\_NOTAIL\_FL). (not used by ext4)
+   * - 0x10000
+     - All directory entry data should be written synchronously (see
+       ``dirsync``) (EXT4\_DIRSYNC\_FL).
+   * - 0x20000
+     - Top of directory hierarchy (EXT4\_TOPDIR\_FL).
+   * - 0x40000
+     - This is a huge file (EXT4\_HUGE\_FILE\_FL).
+   * - 0x80000
+     - Inode uses extents (EXT4\_EXTENTS\_FL).
+   * - 0x200000
+     - Inode stores a large extended attribute value in its data blocks
+       (EXT4\_EA\_INODE\_FL).
+   * - 0x400000
+     - This file has blocks allocated past EOF (EXT4\_EOFBLOCKS\_FL).
+       (deprecated)
+   * - 0x01000000
+     - Inode is a snapshot (``EXT4_SNAPFILE_FL``). (not in mainline)
+   * - 0x04000000
+     - Snapshot is being deleted (``EXT4_SNAPFILE_DELETED_FL``). (not in
+       mainline)
+   * - 0x08000000
+     - Snapshot shrink has completed (``EXT4_SNAPFILE_SHRUNK_FL``). (not in
+       mainline)
+   * - 0x10000000
+     - Inode has inline data (EXT4\_INLINE\_DATA\_FL).
+   * - 0x20000000
+     - Create children with the same project ID (EXT4\_PROJINHERIT\_FL).
+   * - 0x80000000
+     - Reserved for ext4 library (EXT4\_RESERVED\_FL).
+   * -
+     - Aggregate flags:
+   * - 0x4BDFFF
+     - User-visible flags.
+   * - 0x4B80FF
+     - User-modifiable flags. Note that while EXT4\_JOURNAL\_DATA\_FL and
+       EXT4\_EXTENTS\_FL can be set with setattr, they are not in the kernel's
+       EXT4\_FL\_USER\_MODIFIABLE mask, since it needs to handle the setting of
+       these flags in a special manner and they are masked out of the set of
+       flags that are saved directly to i\_flags.
+
+.. _i_osd1:
+
+The ``osd1`` field has multiple meanings depending on the creator:
+
+Linux:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - l\_i\_version
+     - Inode version. However, if the EA\_INODE inode flag is set, this inode
+       stores an extended attribute value and this field contains the upper 32
+       bits of the attribute value's reference count.
+
+Hurd:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - h\_i\_translator
+     - ??
+
+Masix:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - m\_i\_reserved
+     - ??
+
+.. _i_osd2:
+
+The ``osd2`` field has multiple meanings depending on the filesystem creator:
+
+Linux:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le16
+     - l\_i\_blocks\_high
+     - Upper 16-bits of the block count. Please see the note attached to
+       i\_blocks\_lo.
+   * - 0x2
+     - \_\_le16
+     - l\_i\_file\_acl\_high
+     - Upper 16-bits of the extended attribute block (historically, the file
+       ACL location). See the Extended Attributes section below.
+   * - 0x4
+     - \_\_le16
+     - l\_i\_uid\_high
+     - Upper 16-bits of the Owner UID.
+   * - 0x6
+     - \_\_le16
+     - l\_i\_gid\_high
+     - Upper 16-bits of the GID.
+   * - 0x8
+     - \_\_le16
+     - l\_i\_checksum\_lo
+     - Lower 16-bits of the inode checksum.
+   * - 0xA
+     - \_\_le16
+     - l\_i\_reserved
+     - Unused.
+
+Hurd:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le16
+     - h\_i\_reserved1
+     - ??
+   * - 0x2
+     - \_\_u16
+     - h\_i\_mode\_high
+     - Upper 16-bits of the file mode.
+   * - 0x4
+     - \_\_le16
+     - h\_i\_uid\_high
+     - Upper 16-bits of the Owner UID.
+   * - 0x6
+     - \_\_le16
+     - h\_i\_gid\_high
+     - Upper 16-bits of the GID.
+   * - 0x8
+     - \_\_u32
+     - h\_i\_author
+     - Author code?
+
+Masix:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le16
+     - h\_i\_reserved1
+     - ??
+   * - 0x2
+     - \_\_u16
+     - m\_i\_file\_acl\_high
+     - Upper 16-bits of the extended attribute block (historically, the file
+       ACL location).
+   * - 0x4
+     - \_\_u32
+     - m\_i\_reserved2[2]
+     - ??
+
+Inode Size
+~~~~~~~~~~
+
+In ext2 and ext3, the inode structure size was fixed at 128 bytes
+(``EXT2_GOOD_OLD_INODE_SIZE``) and each inode had a disk record size of
+128 bytes. Starting with ext4, it is possible to allocate a larger
+on-disk inode at format time for all inodes in the filesystem to provide
+space beyond the end of the original ext2 inode. The on-disk inode
+record size is recorded in the superblock as ``s_inode_size``. The
+number of bytes actually used by struct ext4\_inode beyond the original
+128-byte ext2 inode is recorded in the ``i_extra_isize`` field for each
+inode, which allows struct ext4\_inode to grow for a new kernel without
+having to upgrade all of the on-disk inodes. Access to fields beyond
+EXT2\_GOOD\_OLD\_INODE\_SIZE should be verified to be within
+``i_extra_isize``. By default, ext4 inode records are 256 bytes, and (as
+of October 2013) the inode structure is 156 bytes
+(``i_extra_isize = 28``). The extra space between the end of the inode
+structure and the end of the inode record can be used to store extended
+attributes. Each inode record can be as large as the filesystem block
+size, though this is not terribly efficient.
+
+Finding an Inode
+~~~~~~~~~~~~~~~~
+
+Each block group contains ``sb->s_inodes_per_group`` inodes. Because
+inode 0 is defined not to exist, this formula can be used to find the
+block group that an inode lives in:
+``bg = (inode_num - 1) / sb->s_inodes_per_group``. The particular inode
+can be found within the block group's inode table at
+``index = (inode_num - 1) % sb->s_inodes_per_group``. To get the byte
+address within the inode table, use
+``offset = index * sb->s_inode_size``.
+
+Inode Timestamps
+~~~~~~~~~~~~~~~~
+
+Four timestamps are recorded in the lower 128 bytes of the inode
+structure -- inode change time (ctime), access time (atime), data
+modification time (mtime), and deletion time (dtime). The four fields
+are 32-bit signed integers that represent seconds since the Unix epoch
+(1970-01-01 00:00:00 GMT), which means that the fields will overflow in
+January 2038. For inodes that are not linked from any directory but are
+still open (orphan inodes), the dtime field is overloaded for use with
+the orphan list. The superblock field ``s_last_orphan`` points to the
+first inode in the orphan list; dtime is then the number of the next
+orphaned inode, or zero if there are no more orphans.
+
+If the inode structure size ``sb->s_inode_size`` is larger than 128
+bytes and the ``i_inode_extra`` field is large enough to encompass the
+respective ``i_[cma]time_extra`` field, the ctime, atime, and mtime
+inode fields are widened to 64 bits. Within this “extra” 32-bit field,
+the lower two bits are used to extend the 32-bit seconds field to be 34
+bit wide; the upper 30 bits are used to provide nanosecond timestamp
+accuracy. Therefore, timestamps should not overflow until May 2446.
+dtime was not widened. There is also a fifth timestamp to record inode
+creation time (crtime); this field is 64-bits wide and decoded in the
+same manner as 64-bit [cma]time. Neither crtime nor dtime are accessible
+through the regular stat() interface, though debugfs will report them.
+
+We use the 32-bit signed time value plus (2^32 \* (extra epoch bits)).
+In other words:
+
+.. list-table::
+   :widths: 20 20 20 20 20
+   :header-rows: 1
+
+   * - Extra epoch bits
+     - MSB of 32-bit time
+     - Adjustment for signed 32-bit to 64-bit tv\_sec
+     - Decoded 64-bit tv\_sec
+     - valid time range
+   * - 0 0
+     - 1
+     - 0
+     - ``-0x80000000 - -0x00000001``
+     - 1901-12-13 to 1969-12-31
+   * - 0 0
+     - 0
+     - 0
+     - ``0x000000000 - 0x07fffffff``
+     - 1970-01-01 to 2038-01-19
+   * - 0 1
+     - 1
+     - 0x100000000
+     - ``0x080000000 - 0x0ffffffff``
+     - 2038-01-19 to 2106-02-07
+   * - 0 1
+     - 0
+     - 0x100000000
+     - ``0x100000000 - 0x17fffffff``
+     - 2106-02-07 to 2174-02-25
+   * - 1 0
+     - 1
+     - 0x200000000
+     - ``0x180000000 - 0x1ffffffff``
+     - 2174-02-25 to 2242-03-16
+   * - 1 0
+     - 0
+     - 0x200000000
+     - ``0x200000000 - 0x27fffffff``
+     - 2242-03-16 to 2310-04-04
+   * - 1 1
+     - 1
+     - 0x300000000
+     - ``0x280000000 - 0x2ffffffff``
+     - 2310-04-04 to 2378-04-22
+   * - 1 1
+     - 0
+     - 0x300000000
+     - ``0x300000000 - 0x37fffffff``
+     - 2378-04-22 to 2446-05-10
+
+This is a somewhat odd encoding since there are effectively seven times
+as many positive values as negative values. There have also been
+long-standing bugs decoding and encoding dates beyond 2038, which don't
+seem to be fixed as of kernel 3.12 and e2fsprogs 1.42.8. 64-bit kernels
+incorrectly use the extra epoch bits 1,1 for dates between 1901 and
+1970. At some point the kernel will be fixed and e2fsck will fix this
+situation, assuming that it is run before 2310.
diff --git a/Documentation/filesystems/ext4/ondisk/journal.rst b/Documentation/filesystems/ext4/ondisk/journal.rst
new file mode 100644
index 000000000..e7031af86
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/journal.rst
@@ -0,0 +1,611 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Journal (jbd2)
+--------------
+
+Introduced in ext3, the ext4 filesystem employs a journal to protect the
+filesystem against corruption in the case of a system crash. A small
+continuous region of disk (default 128MiB) is reserved inside the
+filesystem as a place to land “important” data writes on-disk as quickly
+as possible. Once the important data transaction is fully written to the
+disk and flushed from the disk write cache, a record of the data being
+committed is also written to the journal. At some later point in time,
+the journal code writes the transactions to their final locations on
+disk (this could involve a lot of seeking or a lot of small
+read-write-erases) before erasing the commit record. Should the system
+crash during the second slow write, the journal can be replayed all the
+way to the latest commit record, guaranteeing the atomicity of whatever
+gets written through the journal to the disk. The effect of this is to
+guarantee that the filesystem does not become stuck midway through a
+metadata update.
+
+For performance reasons, ext4 by default only writes filesystem metadata
+through the journal. This means that file data blocks are /not/
+guaranteed to be in any consistent state after a crash. If this default
+guarantee level (``data=ordered``) is not satisfactory, there is a mount
+option to control journal behavior. If ``data=journal``, all data and
+metadata are written to disk through the journal. This is slower but
+safest. If ``data=writeback``, dirty data blocks are not flushed to the
+disk before the metadata are written to disk through the journal.
+
+The journal inode is typically inode 8. The first 68 bytes of the
+journal inode are replicated in the ext4 superblock. The journal itself
+is normal (but hidden) file within the filesystem. The file usually
+consumes an entire block group, though mke2fs tries to put it in the
+middle of the disk.
+
+All fields in jbd2 are written to disk in big-endian order. This is the
+opposite of ext4.
+
+NOTE: Both ext4 and ocfs2 use jbd2.
+
+The maximum size of a journal embedded in an ext4 filesystem is 2^32
+blocks. jbd2 itself does not seem to care.
+
+Layout
+~~~~~~
+
+Generally speaking, the journal has this format:
+
+.. list-table::
+   :widths: 1 1 78
+   :header-rows: 1
+
+   * - Superblock
+     - descriptor\_block (data\_blocks or revocation\_block) [more data or
+       revocations] commmit\_block
+     - [more transactions...]
+   * - 
+     - One transaction
+     -
+
+Notice that a transaction begins with either a descriptor and some data,
+or a block revocation list. A finished transaction always ends with a
+commit. If there is no commit record (or the checksums don't match), the
+transaction will be discarded during replay.
+
+External Journal
+~~~~~~~~~~~~~~~~
+
+Optionally, an ext4 filesystem can be created with an external journal
+device (as opposed to an internal journal, which uses a reserved inode).
+In this case, on the filesystem device, ``s_journal_inum`` should be
+zero and ``s_journal_uuid`` should be set. On the journal device there
+will be an ext4 super block in the usual place, with a matching UUID.
+The journal superblock will be in the next full block after the
+superblock.
+
+.. list-table::
+   :widths: 1 1 1 1 76
+   :header-rows: 1
+
+   * - 1024 bytes of padding
+     - ext4 Superblock
+     - Journal Superblock
+     - descriptor\_block (data\_blocks or revocation\_block) [more data or
+       revocations] commmit\_block
+     - [more transactions...]
+   * - 
+     -
+     -
+     - One transaction
+     -
+
+Block Header
+~~~~~~~~~~~~
+
+Every block in the journal starts with a common 12-byte header
+``struct journal_header_s``:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Description
+   * - 0x0
+     - \_\_be32
+     - h\_magic
+     - jbd2 magic number, 0xC03B3998.
+   * - 0x4
+     - \_\_be32
+     - h\_blocktype
+     - Description of what this block contains. See the jbd2_blocktype_ table
+       below.
+   * - 0x8
+     - \_\_be32
+     - h\_sequence
+     - The transaction ID that goes with this block.
+
+.. _jbd2_blocktype:
+
+The journal block type can be any one of:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 1
+     - Descriptor. This block precedes a series of data blocks that were
+       written through the journal during a transaction.
+   * - 2
+     - Block commit record. This block signifies the completion of a
+       transaction.
+   * - 3
+     - Journal superblock, v1.
+   * - 4
+     - Journal superblock, v2.
+   * - 5
+     - Block revocation records. This speeds up recovery by enabling the
+       journal to skip writing blocks that were subsequently rewritten.
+
+Super Block
+~~~~~~~~~~~
+
+The super block for the journal is much simpler as compared to ext4's.
+The key data kept within are size of the journal, and where to find the
+start of the log of transactions.
+
+The journal superblock is recorded as ``struct journal_superblock_s``,
+which is 1024 bytes long:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Description
+   * -
+     -
+     -
+     - Static information describing the journal.
+   * - 0x0
+     - journal\_header\_t (12 bytes)
+     - s\_header
+     - Common header identifying this as a superblock.
+   * - 0xC
+     - \_\_be32
+     - s\_blocksize
+     - Journal device block size.
+   * - 0x10
+     - \_\_be32
+     - s\_maxlen
+     - Total number of blocks in this journal.
+   * - 0x14
+     - \_\_be32
+     - s\_first
+     - First block of log information.
+   * -
+     -
+     -
+     - Dynamic information describing the current state of the log.
+   * - 0x18
+     - \_\_be32
+     - s\_sequence
+     - First commit ID expected in log.
+   * - 0x1C
+     - \_\_be32
+     - s\_start
+     - Block number of the start of log. Contrary to the comments, this field
+       being zero does not imply that the journal is clean!
+   * - 0x20
+     - \_\_be32
+     - s\_errno
+     - Error value, as set by jbd2\_journal\_abort().
+   * -
+     -
+     -
+     - The remaining fields are only valid in a v2 superblock.
+   * - 0x24
+     - \_\_be32
+     - s\_feature\_compat;
+     - Compatible feature set. See the table jbd2_compat_ below.
+   * - 0x28
+     - \_\_be32
+     - s\_feature\_incompat
+     - Incompatible feature set. See the table jbd2_incompat_ below.
+   * - 0x2C
+     - \_\_be32
+     - s\_feature\_ro\_compat
+     - Read-only compatible feature set. There aren't any of these currently.
+   * - 0x30
+     - \_\_u8
+     - s\_uuid[16]
+     - 128-bit uuid for journal. This is compared against the copy in the ext4
+       super block at mount time.
+   * - 0x40
+     - \_\_be32
+     - s\_nr\_users
+     - Number of file systems sharing this journal.
+   * - 0x44
+     - \_\_be32
+     - s\_dynsuper
+     - Location of dynamic super block copy. (Not used?)
+   * - 0x48
+     - \_\_be32
+     - s\_max\_transaction
+     - Limit of journal blocks per transaction. (Not used?)
+   * - 0x4C
+     - \_\_be32
+     - s\_max\_trans\_data
+     - Limit of data blocks per transaction. (Not used?)
+   * - 0x50
+     - \_\_u8
+     - s\_checksum\_type
+     - Checksum algorithm used for the journal.  See jbd2_checksum_type_ for
+       more info.
+   * - 0x51
+     - \_\_u8[3]
+     - s\_padding2
+     -
+   * - 0x54
+     - \_\_u32
+     - s\_padding[42]
+     -
+   * - 0xFC
+     - \_\_be32
+     - s\_checksum
+     - Checksum of the entire superblock, with this field set to zero.
+   * - 0x100
+     - \_\_u8
+     - s\_users[16\*48]
+     - ids of all file systems sharing the log. e2fsprogs/Linux don't allow
+       shared external journals, but I imagine Lustre (or ocfs2?), which use
+       the jbd2 code, might.
+
+.. _jbd2_compat:
+
+The journal compat features are any combination of the following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x1
+     - Journal maintains checksums on the data blocks.
+       (JBD2\_FEATURE\_COMPAT\_CHECKSUM)
+
+.. _jbd2_incompat:
+
+The journal incompat features are any combination of the following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x1
+     - Journal has block revocation records. (JBD2\_FEATURE\_INCOMPAT\_REVOKE)
+   * - 0x2
+     - Journal can deal with 64-bit block numbers.
+       (JBD2\_FEATURE\_INCOMPAT\_64BIT)
+   * - 0x4
+     - Journal commits asynchronously. (JBD2\_FEATURE\_INCOMPAT\_ASYNC\_COMMIT)
+   * - 0x8
+     - This journal uses v2 of the checksum on-disk format. Each journal
+       metadata block gets its own checksum, and the block tags in the
+       descriptor table contain checksums for each of the data blocks in the
+       journal. (JBD2\_FEATURE\_INCOMPAT\_CSUM\_V2)
+   * - 0x10
+     - This journal uses v3 of the checksum on-disk format. This is the same as
+       v2, but the journal block tag size is fixed regardless of the size of
+       block numbers. (JBD2\_FEATURE\_INCOMPAT\_CSUM\_V3)
+
+.. _jbd2_checksum_type:
+
+Journal checksum type codes are one of the following.  crc32 or crc32c are the
+most likely choices.
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 1
+     - CRC32
+   * - 2
+     - MD5
+   * - 3
+     - SHA1
+   * - 4
+     - CRC32C
+
+Descriptor Block
+~~~~~~~~~~~~~~~~
+
+The descriptor block contains an array of journal block tags that
+describe the final locations of the data blocks that follow in the
+journal. Descriptor blocks are open-coded instead of being completely
+described by a data structure, but here is the block structure anyway.
+Descriptor blocks consume at least 36 bytes, but use a full block:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Descriptor
+   * - 0x0
+     - journal\_header\_t
+     - (open coded)
+     - Common block header.
+   * - 0xC
+     - struct journal\_block\_tag\_s
+     - open coded array[]
+     - Enough tags either to fill up the block or to describe all the data
+       blocks that follow this descriptor block.
+
+Journal block tags have any of the following formats, depending on which
+journal feature and block tag flags are set.
+
+If JBD2\_FEATURE\_INCOMPAT\_CSUM\_V3 is set, the journal block tag is
+defined as ``struct journal_block_tag3_s``, which looks like the
+following. The size is 16 or 32 bytes.
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Descriptor
+   * - 0x0
+     - \_\_be32
+     - t\_blocknr
+     - Lower 32-bits of the location of where the corresponding data block
+       should end up on disk.
+   * - 0x4
+     - \_\_be32
+     - t\_flags
+     - Flags that go with the descriptor. See the table jbd2_tag_flags_ for
+       more info.
+   * - 0x8
+     - \_\_be32
+     - t\_blocknr\_high
+     - Upper 32-bits of the location of where the corresponding data block
+       should end up on disk. This is zero if JBD2\_FEATURE\_INCOMPAT\_64BIT is
+       not enabled.
+   * - 0xC
+     - \_\_be32
+     - t\_checksum
+     - Checksum of the journal UUID, the sequence number, and the data block.
+   * -
+     -
+     -
+     - This field appears to be open coded. It always comes at the end of the
+       tag, after t_checksum. This field is not present if the "same UUID" flag
+       is set.
+   * - 0x8 or 0xC
+     - char
+     - uuid[16]
+     - A UUID to go with this tag. This field appears to be copied from the
+       ``j_uuid`` field in ``struct journal_s``, but only tune2fs touches that
+       field.
+
+.. _jbd2_tag_flags:
+
+The journal tag flags are any combination of the following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x1
+     - On-disk block is escaped. The first four bytes of the data block just
+       happened to match the jbd2 magic number.
+   * - 0x2
+     - This block has the same UUID as previous, therefore the UUID field is
+       omitted.
+   * - 0x4
+     - The data block was deleted by the transaction. (Not used?)
+   * - 0x8
+     - This is the last tag in this descriptor block.
+
+If JBD2\_FEATURE\_INCOMPAT\_CSUM\_V3 is NOT set, the journal block tag
+is defined as ``struct journal_block_tag_s``, which looks like the
+following. The size is 8, 12, 24, or 28 bytes:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Descriptor
+   * - 0x0
+     - \_\_be32
+     - t\_blocknr
+     - Lower 32-bits of the location of where the corresponding data block
+       should end up on disk.
+   * - 0x4
+     - \_\_be16
+     - t\_checksum
+     - Checksum of the journal UUID, the sequence number, and the data block.
+       Note that only the lower 16 bits are stored.
+   * - 0x6
+     - \_\_be16
+     - t\_flags
+     - Flags that go with the descriptor. See the table jbd2_tag_flags_ for
+       more info.
+   * -
+     -
+     -
+     - This next field is only present if the super block indicates support for
+       64-bit block numbers.
+   * - 0x8
+     - \_\_be32
+     - t\_blocknr\_high
+     - Upper 32-bits of the location of where the corresponding data block
+       should end up on disk.
+   * -
+     -
+     -
+     - This field appears to be open coded. It always comes at the end of the
+       tag, after t_flags or t_blocknr_high. This field is not present if the
+       "same UUID" flag is set.
+   * - 0x8 or 0xC
+     - char
+     - uuid[16]
+     - A UUID to go with this tag. This field appears to be copied from the
+       ``j_uuid`` field in ``struct journal_s``, but only tune2fs touches that
+       field.
+
+If JBD2\_FEATURE\_INCOMPAT\_CSUM\_V2 or
+JBD2\_FEATURE\_INCOMPAT\_CSUM\_V3 are set, the end of the block is a
+``struct jbd2_journal_block_tail``, which looks like this:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Descriptor
+   * - 0x0
+     - \_\_be32
+     - t\_checksum
+     - Checksum of the journal UUID + the descriptor block, with this field set
+       to zero.
+
+Data Block
+~~~~~~~~~~
+
+In general, the data blocks being written to disk through the journal
+are written verbatim into the journal file after the descriptor block.
+However, if the first four bytes of the block match the jbd2 magic
+number then those four bytes are replaced with zeroes and the “escaped”
+flag is set in the descriptor block tag.
+
+Revocation Block
+~~~~~~~~~~~~~~~~
+
+A revocation block is used to prevent replay of a block in an earlier
+transaction. This is used to mark blocks that were journalled at one
+time but are no longer journalled. Typically this happens if a metadata
+block is freed and re-allocated as a file data block; in this case, a
+journal replay after the file block was written to disk will cause
+corruption.
+
+**NOTE**: This mechanism is NOT used to express “this journal block is
+superseded by this other journal block”, as the author (djwong)
+mistakenly thought. Any block being added to a transaction will cause
+the removal of all existing revocation records for that block.
+
+Revocation blocks are described in
+``struct jbd2_journal_revoke_header_s``, are at least 16 bytes in
+length, but use a full block:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Description
+   * - 0x0
+     - journal\_header\_t
+     - r\_header
+     - Common block header.
+   * - 0xC
+     - \_\_be32
+     - r\_count
+     - Number of bytes used in this block.
+   * - 0x10
+     - \_\_be32 or \_\_be64
+     - blocks[0]
+     - Blocks to revoke.
+
+After r\_count is a linear array of block numbers that are effectively
+revoked by this transaction. The size of each block number is 8 bytes if
+the superblock advertises 64-bit block number support, or 4 bytes
+otherwise.
+
+If JBD2\_FEATURE\_INCOMPAT\_CSUM\_V2 or
+JBD2\_FEATURE\_INCOMPAT\_CSUM\_V3 are set, the end of the revocation
+block is a ``struct jbd2_journal_revoke_tail``, which has this format:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Description
+   * - 0x0
+     - \_\_be32
+     - r\_checksum
+     - Checksum of the journal UUID + revocation block
+
+Commit Block
+~~~~~~~~~~~~
+
+The commit block is a sentry that indicates that a transaction has been
+completely written to the journal. Once this commit block reaches the
+journal, the data stored with this transaction can be written to their
+final locations on disk.
+
+The commit block is described by ``struct commit_header``, which is 32
+bytes long (but uses a full block):
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Descriptor
+   * - 0x0
+     - journal\_header\_s
+     - (open coded)
+     - Common block header.
+   * - 0xC
+     - unsigned char
+     - h\_chksum\_type
+     - The type of checksum to use to verify the integrity of the data blocks
+       in the transaction. See jbd2_checksum_type_ for more info.
+   * - 0xD
+     - unsigned char
+     - h\_chksum\_size
+     - The number of bytes used by the checksum. Most likely 4.
+   * - 0xE
+     - unsigned char
+     - h\_padding[2]
+     -
+   * - 0x10
+     - \_\_be32
+     - h\_chksum[JBD2\_CHECKSUM\_BYTES]
+     - 32 bytes of space to store checksums. If
+       JBD2\_FEATURE\_INCOMPAT\_CSUM\_V2 or JBD2\_FEATURE\_INCOMPAT\_CSUM\_V3
+       are set, the first ``__be32`` is the checksum of the journal UUID and
+       the entire commit block, with this field zeroed. If
+       JBD2\_FEATURE\_COMPAT\_CHECKSUM is set, the first ``__be32`` is the
+       crc32 of all the blocks already written to the transaction.
+   * - 0x30
+     - \_\_be64
+     - h\_commit\_sec
+     - The time that the transaction was committed, in seconds since the epoch.
+   * - 0x38
+     - \_\_be32
+     - h\_commit\_nsec
+     - Nanoseconds component of the above timestamp.
+
diff --git a/Documentation/filesystems/ext4/ondisk/mmp.rst b/Documentation/filesystems/ext4/ondisk/mmp.rst
new file mode 100644
index 000000000..b7d7a3137
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/mmp.rst
@@ -0,0 +1,77 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Multiple Mount Protection
+-------------------------
+
+Multiple mount protection (MMP) is a feature that protects the
+filesystem against multiple hosts trying to use the filesystem
+simultaneously. When a filesystem is opened (for mounting, or fsck,
+etc.), the MMP code running on the node (call it node A) checks a
+sequence number. If the sequence number is EXT4\_MMP\_SEQ\_CLEAN, the
+open continues. If the sequence number is EXT4\_MMP\_SEQ\_FSCK, then
+fsck is (hopefully) running, and open fails immediately. Otherwise, the
+open code will wait for twice the specified MMP check interval and check
+the sequence number again. If the sequence number has changed, then the
+filesystem is active on another machine and the open fails. If the MMP
+code passes all of those checks, a new MMP sequence number is generated
+and written to the MMP block, and the mount proceeds.
+
+While the filesystem is live, the kernel sets up a timer to re-check the
+MMP block at the specified MMP check interval. To perform the re-check,
+the MMP sequence number is re-read; if it does not match the in-memory
+MMP sequence number, then another node (node B) has mounted the
+filesystem, and node A remounts the filesystem read-only. If the
+sequence numbers match, the sequence number is incremented both in
+memory and on disk, and the re-check is complete.
+
+The hostname and device filename are written into the MMP block whenever
+an open operation succeeds. The MMP code does not use these values; they
+are provided purely for informational purposes.
+
+The checksum is calculated against the FS UUID and the MMP structure.
+The MMP structure (``struct mmp_struct``) is as follows:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Type
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - mmp\_magic
+     - Magic number for MMP, 0x004D4D50 (“MMP”).
+   * - 0x4
+     - \_\_le32
+     - mmp\_seq
+     - Sequence number, updated periodically.
+   * - 0x8
+     - \_\_le64
+     - mmp\_time
+     - Time that the MMP block was last updated.
+   * - 0x10
+     - char[64]
+     - mmp\_nodename
+     - Hostname of the node that opened the filesystem.
+   * - 0x50
+     - char[32]
+     - mmp\_bdevname
+     - Block device name of the filesystem.
+   * - 0x70
+     - \_\_le16
+     - mmp\_check\_interval
+     - The MMP re-check interval, in seconds.
+   * - 0x72
+     - \_\_le16
+     - mmp\_pad1
+     - Zero.
+   * - 0x74
+     - \_\_le32[226]
+     - mmp\_pad2
+     - Zero.
+   * - 0x3FC
+     - \_\_le32
+     - mmp\_checksum
+     - Checksum of the MMP block.
diff --git a/Documentation/filesystems/ext4/ondisk/overview.rst b/Documentation/filesystems/ext4/ondisk/overview.rst
new file mode 100644
index 000000000..cbab18bab
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/overview.rst
@@ -0,0 +1,26 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+High Level Design
+=================
+
+An ext4 file system is split into a series of block groups. To reduce
+performance difficulties due to fragmentation, the block allocator tries
+very hard to keep each file's blocks within the same group, thereby
+reducing seek times. The size of a block group is specified in
+``sb.s_blocks_per_group`` blocks, though it can also calculated as 8 \*
+``block_size_in_bytes``. With the default block size of 4KiB, each group
+will contain 32,768 blocks, for a length of 128MiB. The number of block
+groups is the size of the device divided by the size of a block group.
+
+All fields in ext4 are written to disk in little-endian order. HOWEVER,
+all fields in jbd2 (the journal) are written to disk in big-endian
+order.
+
+.. include:: blocks.rst
+.. include:: blockgroup.rst
+.. include:: special_inodes.rst
+.. include:: allocators.rst
+.. include:: checksums.rst
+.. include:: bigalloc.rst
+.. include:: inlinedata.rst
+.. include:: eainode.rst
diff --git a/Documentation/filesystems/ext4/ondisk/special_inodes.rst b/Documentation/filesystems/ext4/ondisk/special_inodes.rst
new file mode 100644
index 000000000..a82f70c9b
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/special_inodes.rst
@@ -0,0 +1,38 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Special inodes
+--------------
+
+ext4 reserves some inode for special features, as follows:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - inode Number
+     - Purpose
+   * - 0
+     - Doesn't exist; there is no inode 0.
+   * - 1
+     - List of defective blocks.
+   * - 2
+     - Root directory.
+   * - 3
+     - User quota.
+   * - 4
+     - Group quota.
+   * - 5
+     - Boot loader.
+   * - 6
+     - Undelete directory.
+   * - 7
+     - Reserved group descriptors inode. (“resize inode”)
+   * - 8
+     - Journal inode.
+   * - 9
+     - The “exclude” inode, for snapshots(?)
+   * - 10
+     - Replica inode, used for some non-upstream feature?
+   * - 11
+     - Traditional first non-reserved inode. Usually this is the lost+found directory. See s\_first\_ino in the superblock.
+
diff --git a/Documentation/filesystems/ext4/ondisk/super.rst b/Documentation/filesystems/ext4/ondisk/super.rst
new file mode 100644
index 000000000..5f81dd87e
--- /dev/null
+++ b/Documentation/filesystems/ext4/ondisk/super.rst
@@ -0,0 +1,801 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+Super Block
+-----------
+
+The superblock records various information about the enclosing
+filesystem, such as block counts, inode counts, supported features,
+maintenance information, and more.
+
+If the sparse\_super feature flag is set, redundant copies of the
+superblock and group descriptors are kept only in the groups whose group
+number is either 0 or a power of 3, 5, or 7. If the flag is not set,
+redundant copies are kept in all groups.
+
+The superblock checksum is calculated against the superblock structure,
+which includes the FS UUID.
+
+The ext4 superblock is laid out as follows in
+``struct ext4_super_block``:
+
+.. list-table::
+   :widths: 1 1 1 77
+   :header-rows: 1
+
+   * - Offset
+     - Size
+     - Name
+     - Description
+   * - 0x0
+     - \_\_le32
+     - s\_inodes\_count
+     - Total inode count.
+   * - 0x4
+     - \_\_le32
+     - s\_blocks\_count\_lo
+     - Total block count.
+   * - 0x8
+     - \_\_le32
+     - s\_r\_blocks\_count\_lo
+     - This number of blocks can only be allocated by the super-user.
+   * - 0xC
+     - \_\_le32
+     - s\_free\_blocks\_count\_lo
+     - Free block count.
+   * - 0x10
+     - \_\_le32
+     - s\_free\_inodes\_count
+     - Free inode count.
+   * - 0x14
+     - \_\_le32
+     - s\_first\_data\_block
+     - First data block. This must be at least 1 for 1k-block filesystems and
+       is typically 0 for all other block sizes.
+   * - 0x18
+     - \_\_le32
+     - s\_log\_block\_size
+     - Block size is 2 ^ (10 + s\_log\_block\_size).
+   * - 0x1C
+     - \_\_le32
+     - s\_log\_cluster\_size
+     - Cluster size is (2 ^ s\_log\_cluster\_size) blocks if bigalloc is
+       enabled. Otherwise s\_log\_cluster\_size must equal s\_log\_block\_size.
+   * - 0x20
+     - \_\_le32
+     - s\_blocks\_per\_group
+     - Blocks per group.
+   * - 0x24
+     - \_\_le32
+     - s\_clusters\_per\_group
+     - Clusters per group, if bigalloc is enabled. Otherwise
+       s\_clusters\_per\_group must equal s\_blocks\_per\_group.
+   * - 0x28
+     - \_\_le32
+     - s\_inodes\_per\_group
+     - Inodes per group.
+   * - 0x2C
+     - \_\_le32
+     - s\_mtime
+     - Mount time, in seconds since the epoch.
+   * - 0x30
+     - \_\_le32
+     - s\_wtime
+     - Write time, in seconds since the epoch.
+   * - 0x34
+     - \_\_le16
+     - s\_mnt\_count
+     - Number of mounts since the last fsck.
+   * - 0x36
+     - \_\_le16
+     - s\_max\_mnt\_count
+     - Number of mounts beyond which a fsck is needed.
+   * - 0x38
+     - \_\_le16
+     - s\_magic
+     - Magic signature, 0xEF53
+   * - 0x3A
+     - \_\_le16
+     - s\_state
+     - File system state. See super_state_ for more info.
+   * - 0x3C
+     - \_\_le16
+     - s\_errors
+     - Behaviour when detecting errors. See super_errors_ for more info.
+   * - 0x3E
+     - \_\_le16
+     - s\_minor\_rev\_level
+     - Minor revision level.
+   * - 0x40
+     - \_\_le32
+     - s\_lastcheck
+     - Time of last check, in seconds since the epoch.
+   * - 0x44
+     - \_\_le32
+     - s\_checkinterval
+     - Maximum time between checks, in seconds.
+   * - 0x48
+     - \_\_le32
+     - s\_creator\_os
+     - Creator OS. See the table super_creator_ for more info.
+   * - 0x4C
+     - \_\_le32
+     - s\_rev\_level
+     - Revision level. See the table super_revision_ for more info.
+   * - 0x50
+     - \_\_le16
+     - s\_def\_resuid
+     - Default uid for reserved blocks.
+   * - 0x52
+     - \_\_le16
+     - s\_def\_resgid
+     - Default gid for reserved blocks.
+   * -
+     -
+     -
+     - These fields are for EXT4_DYNAMIC_REV superblocks only.
+       
+       Note: the difference between the compatible feature set and the
+       incompatible feature set is that if there is a bit set in the
+       incompatible feature set that the kernel doesn't know about, it should
+       refuse to mount the filesystem.
+       
+       e2fsck's requirements are more strict; if it doesn't know
+       about a feature in either the compatible or incompatible feature set, it
+       must abort and not try to meddle with things it doesn't understand...
+   * - 0x54
+     - \_\_le32
+     - s\_first\_ino
+     - First non-reserved inode.
+   * - 0x58
+     - \_\_le16
+     - s\_inode\_size
+     - Size of inode structure, in bytes.
+   * - 0x5A
+     - \_\_le16
+     - s\_block\_group\_nr
+     - Block group # of this superblock.
+   * - 0x5C
+     - \_\_le32
+     - s\_feature\_compat
+     - Compatible feature set flags. Kernel can still read/write this fs even
+       if it doesn't understand a flag; fsck should not do that. See the
+       super_compat_ table for more info.
+   * - 0x60
+     - \_\_le32
+     - s\_feature\_incompat
+     - Incompatible feature set. If the kernel or fsck doesn't understand one
+       of these bits, it should stop. See the super_incompat_ table for more
+       info.
+   * - 0x64
+     - \_\_le32
+     - s\_feature\_ro\_compat
+     - Readonly-compatible feature set. If the kernel doesn't understand one of
+       these bits, it can still mount read-only. See the super_rocompat_ table
+       for more info.
+   * - 0x68
+     - \_\_u8
+     - s\_uuid[16]
+     - 128-bit UUID for volume.
+   * - 0x78
+     - char
+     - s\_volume\_name[16]
+     - Volume label.
+   * - 0x88
+     - char
+     - s\_last\_mounted[64]
+     - Directory where filesystem was last mounted.
+   * - 0xC8
+     - \_\_le32
+     - s\_algorithm\_usage\_bitmap
+     - For compression (Not used in e2fsprogs/Linux)
+   * -
+     -
+     -
+     - Performance hints.  Directory preallocation should only happen if the
+       EXT4_FEATURE_COMPAT_DIR_PREALLOC flag is on.
+   * - 0xCC
+     - \_\_u8
+     - s\_prealloc\_blocks
+     - #. of blocks to try to preallocate for ... files? (Not used in
+       e2fsprogs/Linux)
+   * - 0xCD
+     - \_\_u8
+     - s\_prealloc\_dir\_blocks
+     - #. of blocks to preallocate for directories. (Not used in
+       e2fsprogs/Linux)
+   * - 0xCE
+     - \_\_le16
+     - s\_reserved\_gdt\_blocks
+     - Number of reserved GDT entries for future filesystem expansion.
+   * -
+     -
+     -
+     - Journalling support is valid only if EXT4_FEATURE_COMPAT_HAS_JOURNAL is
+       set.
+   * - 0xD0
+     - \_\_u8
+     - s\_journal\_uuid[16]
+     - UUID of journal superblock
+   * - 0xE0
+     - \_\_le32
+     - s\_journal\_inum
+     - inode number of journal file.
+   * - 0xE4
+     - \_\_le32
+     - s\_journal\_dev
+     - Device number of journal file, if the external journal feature flag is
+       set.
+   * - 0xE8
+     - \_\_le32
+     - s\_last\_orphan
+     - Start of list of orphaned inodes to delete.
+   * - 0xEC
+     - \_\_le32
+     - s\_hash\_seed[4]
+     - HTREE hash seed.
+   * - 0xFC
+     - \_\_u8
+     - s\_def\_hash\_version
+     - Default hash algorithm to use for directory hashes. See super_def_hash_
+       for more info.
+   * - 0xFD
+     - \_\_u8
+     - s\_jnl\_backup\_type
+     - If this value is 0 or EXT3\_JNL\_BACKUP\_BLOCKS (1), then the
+       ``s_jnl_blocks`` field contains a duplicate copy of the inode's
+       ``i_block[]`` array and ``i_size``.
+   * - 0xFE
+     - \_\_le16
+     - s\_desc\_size
+     - Size of group descriptors, in bytes, if the 64bit incompat feature flag
+       is set.
+   * - 0x100
+     - \_\_le32
+     - s\_default\_mount\_opts
+     - Default mount options. See the super_mountopts_ table for more info.
+   * - 0x104
+     - \_\_le32
+     - s\_first\_meta\_bg
+     - First metablock block group, if the meta\_bg feature is enabled.
+   * - 0x108
+     - \_\_le32
+     - s\_mkfs\_time
+     - When the filesystem was created, in seconds since the epoch.
+   * - 0x10C
+     - \_\_le32
+     - s\_jnl\_blocks[17]
+     - Backup copy of the journal inode's ``i_block[]`` array in the first 15
+       elements and i\_size\_high and i\_size in the 16th and 17th elements,
+       respectively.
+   * -
+     -
+     -
+     - 64bit support is valid only if EXT4_FEATURE_COMPAT_64BIT is set.
+   * - 0x150
+     - \_\_le32
+     - s\_blocks\_count\_hi
+     - High 32-bits of the block count.
+   * - 0x154
+     - \_\_le32
+     - s\_r\_blocks\_count\_hi
+     - High 32-bits of the reserved block count.
+   * - 0x158
+     - \_\_le32
+     - s\_free\_blocks\_count\_hi
+     - High 32-bits of the free block count.
+   * - 0x15C
+     - \_\_le16
+     - s\_min\_extra\_isize
+     - All inodes have at least # bytes.
+   * - 0x15E
+     - \_\_le16
+     - s\_want\_extra\_isize
+     - New inodes should reserve # bytes.
+   * - 0x160
+     - \_\_le32
+     - s\_flags
+     - Miscellaneous flags. See the super_flags_ table for more info.
+   * - 0x164
+     - \_\_le16
+     - s\_raid\_stride
+     - RAID stride. This is the number of logical blocks read from or written
+       to the disk before moving to the next disk. This affects the placement
+       of filesystem metadata, which will hopefully make RAID storage faster.
+   * - 0x166
+     - \_\_le16
+     - s\_mmp\_interval
+     - #. seconds to wait in multi-mount prevention (MMP) checking. In theory,
+       MMP is a mechanism to record in the superblock which host and device
+       have mounted the filesystem, in order to prevent multiple mounts. This
+       feature does not seem to be implemented...
+   * - 0x168
+     - \_\_le64
+     - s\_mmp\_block
+     - Block # for multi-mount protection data.
+   * - 0x170
+     - \_\_le32
+     - s\_raid\_stripe\_width
+     - RAID stripe width. This is the number of logical blocks read from or
+       written to the disk before coming back to the current disk. This is used
+       by the block allocator to try to reduce the number of read-modify-write
+       operations in a RAID5/6.
+   * - 0x174
+     - \_\_u8
+     - s\_log\_groups\_per\_flex
+     - Size of a flexible block group is 2 ^ ``s_log_groups_per_flex``.
+   * - 0x175
+     - \_\_u8
+     - s\_checksum\_type
+     - Metadata checksum algorithm type. The only valid value is 1 (crc32c).
+   * - 0x176
+     - \_\_le16
+     - s\_reserved\_pad
+     -
+   * - 0x178
+     - \_\_le64
+     - s\_kbytes\_written
+     - Number of KiB written to this filesystem over its lifetime.
+   * - 0x180
+     - \_\_le32
+     - s\_snapshot\_inum
+     - inode number of active snapshot. (Not used in e2fsprogs/Linux.)
+   * - 0x184
+     - \_\_le32
+     - s\_snapshot\_id
+     - Sequential ID of active snapshot. (Not used in e2fsprogs/Linux.)
+   * - 0x188
+     - \_\_le64
+     - s\_snapshot\_r\_blocks\_count
+     - Number of blocks reserved for active snapshot's future use. (Not used in
+       e2fsprogs/Linux.)
+   * - 0x190
+     - \_\_le32
+     - s\_snapshot\_list
+     - inode number of the head of the on-disk snapshot list. (Not used in
+       e2fsprogs/Linux.)
+   * - 0x194
+     - \_\_le32
+     - s\_error\_count
+     - Number of errors seen.
+   * - 0x198
+     - \_\_le32
+     - s\_first\_error\_time
+     - First time an error happened, in seconds since the epoch.
+   * - 0x19C
+     - \_\_le32
+     - s\_first\_error\_ino
+     - inode involved in first error.
+   * - 0x1A0
+     - \_\_le64
+     - s\_first\_error\_block
+     - Number of block involved of first error.
+   * - 0x1A8
+     - \_\_u8
+     - s\_first\_error\_func[32]
+     - Name of function where the error happened.
+   * - 0x1C8
+     - \_\_le32
+     - s\_first\_error\_line
+     - Line number where error happened.
+   * - 0x1CC
+     - \_\_le32
+     - s\_last\_error\_time
+     - Time of most recent error, in seconds since the epoch.
+   * - 0x1D0
+     - \_\_le32
+     - s\_last\_error\_ino
+     - inode involved in most recent error.
+   * - 0x1D4
+     - \_\_le32
+     - s\_last\_error\_line
+     - Line number where most recent error happened.
+   * - 0x1D8
+     - \_\_le64
+     - s\_last\_error\_block
+     - Number of block involved in most recent error.
+   * - 0x1E0
+     - \_\_u8
+     - s\_last\_error\_func[32]
+     - Name of function where the most recent error happened.
+   * - 0x200
+     - \_\_u8
+     - s\_mount\_opts[64]
+     - ASCIIZ string of mount options.
+   * - 0x240
+     - \_\_le32
+     - s\_usr\_quota\_inum
+     - Inode number of user `quota <quota>`__ file.
+   * - 0x244
+     - \_\_le32
+     - s\_grp\_quota\_inum
+     - Inode number of group `quota <quota>`__ file.
+   * - 0x248
+     - \_\_le32
+     - s\_overhead\_blocks
+     - Overhead blocks/clusters in fs. (Huh? This field is always zero, which
+       means that the kernel calculates it dynamically.)
+   * - 0x24C
+     - \_\_le32
+     - s\_backup\_bgs[2]
+     - Block groups containing superblock backups (if sparse\_super2)
+   * - 0x254
+     - \_\_u8
+     - s\_encrypt\_algos[4]
+     - Encryption algorithms in use. There can be up to four algorithms in use
+       at any time; valid algorithm codes are given in the super_encrypt_ table
+       below.
+   * - 0x258
+     - \_\_u8
+     - s\_encrypt\_pw\_salt[16]
+     - Salt for the string2key algorithm for encryption.
+   * - 0x268
+     - \_\_le32
+     - s\_lpf\_ino
+     - Inode number of lost+found
+   * - 0x26C
+     - \_\_le32
+     - s\_prj\_quota\_inum
+     - Inode that tracks project quotas.
+   * - 0x270
+     - \_\_le32
+     - s\_checksum\_seed
+     - Checksum seed used for metadata\_csum calculations. This value is
+       crc32c(~0, $orig\_fs\_uuid).
+   * - 0x274
+     - \_\_u8
+     - s\_wtime_hi
+     - Upper 8 bits of the s_wtime field.
+   * - 0x275
+     - \_\_u8
+     - s\_wtime_hi
+     - Upper 8 bits of the s_mtime field.
+   * - 0x276
+     - \_\_u8
+     - s\_mkfs_time_hi
+     - Upper 8 bits of the s_mkfs_time field.
+   * - 0x277
+     - \_\_u8
+     - s\_lastcheck_hi
+     - Upper 8 bits of the s_lastcheck_hi field.
+   * - 0x278
+     - \_\_u8
+     - s\_first_error_time_hi
+     - Upper 8 bits of the s_first_error_time_hi field.
+   * - 0x279
+     - \_\_u8
+     - s\_last_error_time_hi
+     - Upper 8 bits of the s_last_error_time_hi field.
+   * - 0x27A
+     - \_\_u8[2]
+     - s\_pad
+     - Zero padding.
+   * - 0x27C
+     - \_\_le32
+     - s\_reserved[96]
+     - Padding to the end of the block.
+   * - 0x3FC
+     - \_\_le32
+     - s\_checksum
+     - Superblock checksum.
+
+.. _super_state:
+
+The superblock state is some combination of the following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x0001
+     - Cleanly umounted
+   * - 0x0002
+     - Errors detected
+   * - 0x0004
+     - Orphans being recovered
+
+.. _super_errors:
+
+The superblock error policy is one of the following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 1
+     - Continue
+   * - 2
+     - Remount read-only
+   * - 3
+     - Panic
+
+.. _super_creator:
+
+The filesystem creator is one of the following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0
+     - Linux
+   * - 1
+     - Hurd
+   * - 2
+     - Masix
+   * - 3
+     - FreeBSD
+   * - 4
+     - Lites
+
+.. _super_revision:
+
+The superblock revision is one of the following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0
+     - Original format
+   * - 1
+     - v2 format w/ dynamic inode sizes
+
+Note that ``EXT4_DYNAMIC_REV`` refers to a revision 1 or newer filesystem.
+
+.. _super_compat:
+
+The superblock compatible features field is a combination of any of the
+following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x1
+     - Directory preallocation (COMPAT\_DIR\_PREALLOC).
+   * - 0x2
+     - “imagic inodes”. Not clear from the code what this does
+       (COMPAT\_IMAGIC\_INODES).
+   * - 0x4
+     - Has a journal (COMPAT\_HAS\_JOURNAL).
+   * - 0x8
+     - Supports extended attributes (COMPAT\_EXT\_ATTR).
+   * - 0x10
+     - Has reserved GDT blocks for filesystem expansion
+       (COMPAT\_RESIZE\_INODE). Requires RO\_COMPAT\_SPARSE\_SUPER.
+   * - 0x20
+     - Has directory indices (COMPAT\_DIR\_INDEX).
+   * - 0x40
+     - “Lazy BG”. Not in Linux kernel, seems to have been for uninitialized
+       block groups? (COMPAT\_LAZY\_BG)
+   * - 0x80
+     - “Exclude inode”. Not used. (COMPAT\_EXCLUDE\_INODE).
+   * - 0x100
+     - “Exclude bitmap”. Seems to be used to indicate the presence of
+       snapshot-related exclude bitmaps? Not defined in kernel or used in
+       e2fsprogs (COMPAT\_EXCLUDE\_BITMAP).
+   * - 0x200
+     - Sparse Super Block, v2. If this flag is set, the SB field s\_backup\_bgs
+       points to the two block groups that contain backup superblocks
+       (COMPAT\_SPARSE\_SUPER2).
+
+.. _super_incompat:
+
+The superblock incompatible features field is a combination of any of the
+following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x1
+     - Compression (INCOMPAT\_COMPRESSION).
+   * - 0x2
+     - Directory entries record the file type. See ext4\_dir\_entry\_2 below
+       (INCOMPAT\_FILETYPE).
+   * - 0x4
+     - Filesystem needs recovery (INCOMPAT\_RECOVER).
+   * - 0x8
+     - Filesystem has a separate journal device (INCOMPAT\_JOURNAL\_DEV).
+   * - 0x10
+     - Meta block groups. See the earlier discussion of this feature
+       (INCOMPAT\_META\_BG).
+   * - 0x40
+     - Files in this filesystem use extents (INCOMPAT\_EXTENTS).
+   * - 0x80
+     - Enable a filesystem size of 2^64 blocks (INCOMPAT\_64BIT).
+   * - 0x100
+     - Multiple mount protection. Not implemented (INCOMPAT\_MMP).
+   * - 0x200
+     - Flexible block groups. See the earlier discussion of this feature
+       (INCOMPAT\_FLEX\_BG).
+   * - 0x400
+     - Inodes can be used to store large extended attribute values
+       (INCOMPAT\_EA\_INODE).
+   * - 0x1000
+     - Data in directory entry (INCOMPAT\_DIRDATA). (Not implemented?)
+   * - 0x2000
+     - Metadata checksum seed is stored in the superblock. This feature enables
+       the administrator to change the UUID of a metadata\_csum filesystem
+       while the filesystem is mounted; without it, the checksum definition
+       requires all metadata blocks to be rewritten (INCOMPAT\_CSUM\_SEED).
+   * - 0x4000
+     - Large directory >2GB or 3-level htree (INCOMPAT\_LARGEDIR). Prior to
+       this feature, directories could not be larger than 4GiB and could not
+       have an htree more than 2 levels deep. If this feature is enabled,
+       directories can be larger than 4GiB and have a maximum htree depth of 3.
+   * - 0x8000
+     - Data in inode (INCOMPAT\_INLINE\_DATA).
+   * - 0x10000
+     - Encrypted inodes are present on the filesystem. (INCOMPAT\_ENCRYPT).
+
+.. _super_rocompat:
+
+The superblock read-only compatible features field is a combination of any of
+the following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x1
+     - Sparse superblocks. See the earlier discussion of this feature
+       (RO\_COMPAT\_SPARSE\_SUPER).
+   * - 0x2
+     - This filesystem has been used to store a file greater than 2GiB
+       (RO\_COMPAT\_LARGE\_FILE).
+   * - 0x4
+     - Not used in kernel or e2fsprogs (RO\_COMPAT\_BTREE\_DIR).
+   * - 0x8
+     - This filesystem has files whose sizes are represented in units of
+       logical blocks, not 512-byte sectors. This implies a very large file
+       indeed! (RO\_COMPAT\_HUGE\_FILE)
+   * - 0x10
+     - Group descriptors have checksums. In addition to detecting corruption,
+       this is useful for lazy formatting with uninitialized groups
+       (RO\_COMPAT\_GDT\_CSUM).
+   * - 0x20
+     - Indicates that the old ext3 32,000 subdirectory limit no longer applies
+       (RO\_COMPAT\_DIR\_NLINK). A directory's i\_links\_count will be set to 1
+       if it is incremented past 64,999.
+   * - 0x40
+     - Indicates that large inodes exist on this filesystem
+       (RO\_COMPAT\_EXTRA\_ISIZE).
+   * - 0x80
+     - This filesystem has a snapshot (RO\_COMPAT\_HAS\_SNAPSHOT).
+   * - 0x100
+     - `Quota <Quota>`__ (RO\_COMPAT\_QUOTA).
+   * - 0x200
+     - This filesystem supports “bigalloc”, which means that file extents are
+       tracked in units of clusters (of blocks) instead of blocks
+       (RO\_COMPAT\_BIGALLOC).
+   * - 0x400
+     - This filesystem supports metadata checksumming.
+       (RO\_COMPAT\_METADATA\_CSUM; implies RO\_COMPAT\_GDT\_CSUM, though
+       GDT\_CSUM must not be set)
+   * - 0x800
+     - Filesystem supports replicas. This feature is neither in the kernel nor
+       e2fsprogs. (RO\_COMPAT\_REPLICA)
+   * - 0x1000
+     - Read-only filesystem image; the kernel will not mount this image
+       read-write and most tools will refuse to write to the image.
+       (RO\_COMPAT\_READONLY)
+   * - 0x2000
+     - Filesystem tracks project quotas. (RO\_COMPAT\_PROJECT)
+
+.. _super_def_hash:
+
+The ``s_def_hash_version`` field is one of the following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x0
+     - Legacy.
+   * - 0x1
+     - Half MD4.
+   * - 0x2
+     - Tea.
+   * - 0x3
+     - Legacy, unsigned.
+   * - 0x4
+     - Half MD4, unsigned.
+   * - 0x5
+     - Tea, unsigned.
+
+.. _super_mountopts:
+
+The ``s_default_mount_opts`` field is any combination of the following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x0001
+     - Print debugging info upon (re)mount. (EXT4\_DEFM\_DEBUG)
+   * - 0x0002
+     - New files take the gid of the containing directory (instead of the fsgid
+       of the current process). (EXT4\_DEFM\_BSDGROUPS)
+   * - 0x0004
+     - Support userspace-provided extended attributes. (EXT4\_DEFM\_XATTR\_USER)
+   * - 0x0008
+     - Support POSIX access control lists (ACLs). (EXT4\_DEFM\_ACL)
+   * - 0x0010
+     - Do not support 32-bit UIDs. (EXT4\_DEFM\_UID16)
+   * - 0x0020
+     - All data and metadata are commited to the journal.
+       (EXT4\_DEFM\_JMODE\_DATA)
+   * - 0x0040
+     - All data are flushed to the disk before metadata are committed to the
+       journal. (EXT4\_DEFM\_JMODE\_ORDERED)
+   * - 0x0060
+     - Data ordering is not preserved; data may be written after the metadata
+       has been written. (EXT4\_DEFM\_JMODE\_WBACK)
+   * - 0x0100
+     - Disable write flushes. (EXT4\_DEFM\_NOBARRIER)
+   * - 0x0200
+     - Track which blocks in a filesystem are metadata and therefore should not
+       be used as data blocks. This option will be enabled by default on 3.18,
+       hopefully. (EXT4\_DEFM\_BLOCK\_VALIDITY)
+   * - 0x0400
+     - Enable DISCARD support, where the storage device is told about blocks
+       becoming unused. (EXT4\_DEFM\_DISCARD)
+   * - 0x0800
+     - Disable delayed allocation. (EXT4\_DEFM\_NODELALLOC)
+
+.. _super_flags:
+
+The ``s_flags`` field is any combination of the following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0x0001
+     - Signed directory hash in use.
+   * - 0x0002
+     - Unsigned directory hash in use.
+   * - 0x0004
+     - To test development code.
+
+.. _super_encrypt:
+
+The ``s_encrypt_algos`` list can contain any of the following:
+
+.. list-table::
+   :widths: 1 79
+   :header-rows: 1
+
+   * - Value
+     - Description
+   * - 0
+     - Invalid algorithm (ENCRYPTION\_MODE\_INVALID).
+   * - 1
+     - 256-bit AES in XTS mode (ENCRYPTION\_MODE\_AES\_256\_XTS).
+   * - 2
+     - 256-bit AES in GCM mode (ENCRYPTION\_MODE\_AES\_256\_GCM).
+   * - 3
+     - 256-bit AES in CBC mode (ENCRYPTION\_MODE\_AES\_256\_CBC).
+
+Total size of the superblock is 1024 bytes.
diff --git a/Documentation/filesystems/f2fs.txt b/Documentation/filesystems/f2fs.txt
new file mode 100644
index 000000000..e5edd2968
--- /dev/null
+++ b/Documentation/filesystems/f2fs.txt
@@ -0,0 +1,710 @@
+================================================================================
+WHAT IS Flash-Friendly File System (F2FS)?
+================================================================================
+
+NAND flash memory-based storage devices, such as SSD, eMMC, and SD cards, have
+been equipped on a variety systems ranging from mobile to server systems. Since
+they are known to have different characteristics from the conventional rotating
+disks, a file system, an upper layer to the storage device, should adapt to the
+changes from the sketch in the design level.
+
+F2FS is a file system exploiting NAND flash memory-based storage devices, which
+is based on Log-structured File System (LFS). The design has been focused on
+addressing the fundamental issues in LFS, which are snowball effect of wandering
+tree and high cleaning overhead.
+
+Since a NAND flash memory-based storage device shows different characteristic
+according to its internal geometry or flash memory management scheme, namely FTL,
+F2FS and its tools support various parameters not only for configuring on-disk
+layout, but also for selecting allocation and cleaning algorithms.
+
+The following git tree provides the file system formatting tool (mkfs.f2fs),
+a consistency checking tool (fsck.f2fs), and a debugging tool (dump.f2fs).
+>> git://git.kernel.org/pub/scm/linux/kernel/git/jaegeuk/f2fs-tools.git
+
+For reporting bugs and sending patches, please use the following mailing list:
+>> linux-f2fs-devel@lists.sourceforge.net
+
+================================================================================
+BACKGROUND AND DESIGN ISSUES
+================================================================================
+
+Log-structured File System (LFS)
+--------------------------------
+"A log-structured file system writes all modifications to disk sequentially in
+a log-like structure, thereby speeding up  both file writing and crash recovery.
+The log is the only structure on disk; it contains indexing information so that
+files can be read back from the log efficiently. In order to maintain large free
+areas on disk for fast writing, we divide  the log into segments and use a
+segment cleaner to compress the live information from heavily fragmented
+segments." from Rosenblum, M. and Ousterhout, J. K., 1992, "The design and
+implementation of a log-structured file system", ACM Trans. Computer Systems
+10, 1, 26–52.
+
+Wandering Tree Problem
+----------------------
+In LFS, when a file data is updated and written to the end of log, its direct
+pointer block is updated due to the changed location. Then the indirect pointer
+block is also updated due to the direct pointer block update. In this manner,
+the upper index structures such as inode, inode map, and checkpoint block are
+also updated recursively. This problem is called as wandering tree problem [1],
+and in order to enhance the performance, it should eliminate or relax the update
+propagation as much as possible.
+
+[1] Bityutskiy, A. 2005. JFFS3 design issues. http://www.linux-mtd.infradead.org/
+
+Cleaning Overhead
+-----------------
+Since LFS is based on out-of-place writes, it produces so many obsolete blocks
+scattered across the whole storage. In order to serve new empty log space, it
+needs to reclaim these obsolete blocks seamlessly to users. This job is called
+as a cleaning process.
+
+The process consists of three operations as follows.
+1. A victim segment is selected through referencing segment usage table.
+2. It loads parent index structures of all the data in the victim identified by
+   segment summary blocks.
+3. It checks the cross-reference between the data and its parent index structure.
+4. It moves valid data selectively.
+
+This cleaning job may cause unexpected long delays, so the most important goal
+is to hide the latencies to users. And also definitely, it should reduce the
+amount of valid data to be moved, and move them quickly as well.
+
+================================================================================
+KEY FEATURES
+================================================================================
+
+Flash Awareness
+---------------
+- Enlarge the random write area for better performance, but provide the high
+  spatial locality
+- Align FS data structures to the operational units in FTL as best efforts
+
+Wandering Tree Problem
+----------------------
+- Use a term, “node”, that represents inodes as well as various pointer blocks
+- Introduce Node Address Table (NAT) containing the locations of all the “node”
+  blocks; this will cut off the update propagation.
+
+Cleaning Overhead
+-----------------
+- Support a background cleaning process
+- Support greedy and cost-benefit algorithms for victim selection policies
+- Support multi-head logs for static/dynamic hot and cold data separation
+- Introduce adaptive logging for efficient block allocation
+
+================================================================================
+MOUNT OPTIONS
+================================================================================
+
+background_gc=%s       Turn on/off cleaning operations, namely garbage
+                       collection, triggered in background when I/O subsystem is
+                       idle. If background_gc=on, it will turn on the garbage
+                       collection and if background_gc=off, garbage collection
+                       will be turned off. If background_gc=sync, it will turn
+                       on synchronous garbage collection running in background.
+                       Default value for this option is on. So garbage
+                       collection is on by default.
+disable_roll_forward   Disable the roll-forward recovery routine
+norecovery             Disable the roll-forward recovery routine, mounted read-
+                       only (i.e., -o ro,disable_roll_forward)
+discard/nodiscard      Enable/disable real-time discard in f2fs, if discard is
+                       enabled, f2fs will issue discard/TRIM commands when a
+		       segment is cleaned.
+no_heap                Disable heap-style segment allocation which finds free
+                       segments for data from the beginning of main area, while
+		       for node from the end of main area.
+nouser_xattr           Disable Extended User Attributes. Note: xattr is enabled
+                       by default if CONFIG_F2FS_FS_XATTR is selected.
+noacl                  Disable POSIX Access Control List. Note: acl is enabled
+                       by default if CONFIG_F2FS_FS_POSIX_ACL is selected.
+active_logs=%u         Support configuring the number of active logs. In the
+                       current design, f2fs supports only 2, 4, and 6 logs.
+                       Default number is 6.
+disable_ext_identify   Disable the extension list configured by mkfs, so f2fs
+                       does not aware of cold files such as media files.
+inline_xattr           Enable the inline xattrs feature.
+noinline_xattr         Disable the inline xattrs feature.
+inline_data            Enable the inline data feature: New created small(<~3.4k)
+                       files can be written into inode block.
+inline_dentry          Enable the inline dir feature: data in new created
+                       directory entries can be written into inode block. The
+                       space of inode block which is used to store inline
+                       dentries is limited to ~3.4k.
+noinline_dentry        Disable the inline dentry feature.
+flush_merge	       Merge concurrent cache_flush commands as much as possible
+                       to eliminate redundant command issues. If the underlying
+		       device handles the cache_flush command relatively slowly,
+		       recommend to enable this option.
+nobarrier              This option can be used if underlying storage guarantees
+                       its cached data should be written to the novolatile area.
+		       If this option is set, no cache_flush commands are issued
+		       but f2fs still guarantees the write ordering of all the
+		       data writes.
+fastboot               This option is used when a system wants to reduce mount
+                       time as much as possible, even though normal performance
+		       can be sacrificed.
+extent_cache           Enable an extent cache based on rb-tree, it can cache
+                       as many as extent which map between contiguous logical
+                       address and physical address per inode, resulting in
+                       increasing the cache hit ratio. Set by default.
+noextent_cache         Disable an extent cache based on rb-tree explicitly, see
+                       the above extent_cache mount option.
+noinline_data          Disable the inline data feature, inline data feature is
+                       enabled by default.
+data_flush             Enable data flushing before checkpoint in order to
+                       persist data of regular and symlink.
+fault_injection=%d     Enable fault injection in all supported types with
+                       specified injection rate.
+fault_type=%d          Support configuring fault injection type, should be
+                       enabled with fault_injection option, fault type value
+                       is shown below, it supports single or combined type.
+                       Type_Name		Type_Value
+                       FAULT_KMALLOC		0x000000001
+                       FAULT_KVMALLOC		0x000000002
+                       FAULT_PAGE_ALLOC		0x000000004
+                       FAULT_PAGE_GET		0x000000008
+                       FAULT_ALLOC_BIO		0x000000010
+                       FAULT_ALLOC_NID		0x000000020
+                       FAULT_ORPHAN		0x000000040
+                       FAULT_BLOCK		0x000000080
+                       FAULT_DIR_DEPTH		0x000000100
+                       FAULT_EVICT_INODE	0x000000200
+                       FAULT_TRUNCATE		0x000000400
+                       FAULT_IO			0x000000800
+                       FAULT_CHECKPOINT		0x000001000
+                       FAULT_DISCARD		0x000002000
+mode=%s                Control block allocation mode which supports "adaptive"
+                       and "lfs". In "lfs" mode, there should be no random
+                       writes towards main area.
+io_bits=%u             Set the bit size of write IO requests. It should be set
+                       with "mode=lfs".
+usrquota               Enable plain user disk quota accounting.
+grpquota               Enable plain group disk quota accounting.
+prjquota               Enable plain project quota accounting.
+usrjquota=<file>       Appoint specified file and type during mount, so that quota
+grpjquota=<file>       information can be properly updated during recovery flow,
+prjjquota=<file>       <quota file>: must be in root directory;
+jqfmt=<quota type>     <quota type>: [vfsold,vfsv0,vfsv1].
+offusrjquota           Turn off user journelled quota.
+offgrpjquota           Turn off group journelled quota.
+offprjjquota           Turn off project journelled quota.
+quota                  Enable plain user disk quota accounting.
+noquota                Disable all plain disk quota option.
+whint_mode=%s          Control which write hints are passed down to block
+                       layer. This supports "off", "user-based", and
+                       "fs-based".  In "off" mode (default), f2fs does not pass
+                       down hints. In "user-based" mode, f2fs tries to pass
+                       down hints given by users. And in "fs-based" mode, f2fs
+                       passes down hints with its policy.
+alloc_mode=%s          Adjust block allocation policy, which supports "reuse"
+                       and "default".
+fsync_mode=%s          Control the policy of fsync. Currently supports "posix",
+                       "strict", and "nobarrier". In "posix" mode, which is
+                       default, fsync will follow POSIX semantics and does a
+                       light operation to improve the filesystem performance.
+                       In "strict" mode, fsync will be heavy and behaves in line
+                       with xfs, ext4 and btrfs, where xfstest generic/342 will
+                       pass, but the performance will regress. "nobarrier" is
+                       based on "posix", but doesn't issue flush command for
+                       non-atomic files likewise "nobarrier" mount option.
+test_dummy_encryption  Enable dummy encryption, which provides a fake fscrypt
+                       context. The fake fscrypt context is used by xfstests.
+
+================================================================================
+DEBUGFS ENTRIES
+================================================================================
+
+/sys/kernel/debug/f2fs/ contains information about all the partitions mounted as
+f2fs. Each file shows the whole f2fs information.
+
+/sys/kernel/debug/f2fs/status includes:
+ - major file system information managed by f2fs currently
+ - average SIT information about whole segments
+ - current memory footprint consumed by f2fs.
+
+================================================================================
+SYSFS ENTRIES
+================================================================================
+
+Information about mounted f2fs file systems can be found in
+/sys/fs/f2fs.  Each mounted filesystem will have a directory in
+/sys/fs/f2fs based on its device name (i.e., /sys/fs/f2fs/sda).
+The files in each per-device directory are shown in table below.
+
+Files in /sys/fs/f2fs/<devname>
+(see also Documentation/ABI/testing/sysfs-fs-f2fs)
+..............................................................................
+ File                         Content
+
+ gc_max_sleep_time            This tuning parameter controls the maximum sleep
+                              time for the garbage collection thread. Time is
+                              in milliseconds.
+
+ gc_min_sleep_time            This tuning parameter controls the minimum sleep
+                              time for the garbage collection thread. Time is
+                              in milliseconds.
+
+ gc_no_gc_sleep_time          This tuning parameter controls the default sleep
+                              time for the garbage collection thread. Time is
+                              in milliseconds.
+
+ gc_idle                      This parameter controls the selection of victim
+                              policy for garbage collection. Setting gc_idle = 0
+                              (default) will disable this option. Setting
+                              gc_idle = 1 will select the Cost Benefit approach
+                              & setting gc_idle = 2 will select the greedy approach.
+
+ gc_urgent                    This parameter controls triggering background GCs
+                              urgently or not. Setting gc_urgent = 0 [default]
+                              makes back to default behavior, while if it is set
+                              to 1, background thread starts to do GC by given
+                              gc_urgent_sleep_time interval.
+
+ gc_urgent_sleep_time         This parameter controls sleep time for gc_urgent.
+                              500 ms is set by default. See above gc_urgent.
+
+ reclaim_segments             This parameter controls the number of prefree
+                              segments to be reclaimed. If the number of prefree
+			      segments is larger than the number of segments
+			      in the proportion to the percentage over total
+			      volume size, f2fs tries to conduct checkpoint to
+			      reclaim the prefree segments to free segments.
+			      By default, 5% over total # of segments.
+
+ max_small_discards	      This parameter controls the number of discard
+			      commands that consist small blocks less than 2MB.
+			      The candidates to be discarded are cached until
+			      checkpoint is triggered, and issued during the
+			      checkpoint. By default, it is disabled with 0.
+
+ trim_sections                This parameter controls the number of sections
+                              to be trimmed out in batch mode when FITRIM
+                              conducts. 32 sections is set by default.
+
+ ipu_policy                   This parameter controls the policy of in-place
+                              updates in f2fs. There are five policies:
+                               0x01: F2FS_IPU_FORCE, 0x02: F2FS_IPU_SSR,
+                               0x04: F2FS_IPU_UTIL,  0x08: F2FS_IPU_SSR_UTIL,
+                               0x10: F2FS_IPU_FSYNC.
+
+ min_ipu_util                 This parameter controls the threshold to trigger
+                              in-place-updates. The number indicates percentage
+                              of the filesystem utilization, and used by
+                              F2FS_IPU_UTIL and F2FS_IPU_SSR_UTIL policies.
+
+ min_fsync_blocks             This parameter controls the threshold to trigger
+                              in-place-updates when F2FS_IPU_FSYNC mode is set.
+			      The number indicates the number of dirty pages
+			      when fsync needs to flush on its call path. If
+			      the number is less than this value, it triggers
+			      in-place-updates.
+
+ max_victim_search	      This parameter controls the number of trials to
+			      find a victim segment when conducting SSR and
+			      cleaning operations. The default value is 4096
+			      which covers 8GB block address range.
+
+ dir_level                    This parameter controls the directory level to
+			      support large directory. If a directory has a
+			      number of files, it can reduce the file lookup
+			      latency by increasing this dir_level value.
+			      Otherwise, it needs to decrease this value to
+			      reduce the space overhead. The default value is 0.
+
+ ram_thresh                   This parameter controls the memory footprint used
+			      by free nids and cached nat entries. By default,
+			      10 is set, which indicates 10 MB / 1 GB RAM.
+
+================================================================================
+USAGE
+================================================================================
+
+1. Download userland tools and compile them.
+
+2. Skip, if f2fs was compiled statically inside kernel.
+   Otherwise, insert the f2fs.ko module.
+ # insmod f2fs.ko
+
+3. Create a directory trying to mount
+ # mkdir /mnt/f2fs
+
+4. Format the block device, and then mount as f2fs
+ # mkfs.f2fs -l label /dev/block_device
+ # mount -t f2fs /dev/block_device /mnt/f2fs
+
+mkfs.f2fs
+---------
+The mkfs.f2fs is for the use of formatting a partition as the f2fs filesystem,
+which builds a basic on-disk layout.
+
+The options consist of:
+-l [label]   : Give a volume label, up to 512 unicode name.
+-a [0 or 1]  : Split start location of each area for heap-based allocation.
+               1 is set by default, which performs this.
+-o [int]     : Set overprovision ratio in percent over volume size.
+               5 is set by default.
+-s [int]     : Set the number of segments per section.
+               1 is set by default.
+-z [int]     : Set the number of sections per zone.
+               1 is set by default.
+-e [str]     : Set basic extension list. e.g. "mp3,gif,mov"
+-t [0 or 1]  : Disable discard command or not.
+               1 is set by default, which conducts discard.
+
+fsck.f2fs
+---------
+The fsck.f2fs is a tool to check the consistency of an f2fs-formatted
+partition, which examines whether the filesystem metadata and user-made data
+are cross-referenced correctly or not.
+Note that, initial version of the tool does not fix any inconsistency.
+
+The options consist of:
+  -d debug level [default:0]
+
+dump.f2fs
+---------
+The dump.f2fs shows the information of specific inode and dumps SSA and SIT to
+file. Each file is dump_ssa and dump_sit.
+
+The dump.f2fs is used to debug on-disk data structures of the f2fs filesystem.
+It shows on-disk inode information recognized by a given inode number, and is
+able to dump all the SSA and SIT entries into predefined files, ./dump_ssa and
+./dump_sit respectively.
+
+The options consist of:
+  -d debug level [default:0]
+  -i inode no (hex)
+  -s [SIT dump segno from #1~#2 (decimal), for all 0~-1]
+  -a [SSA dump segno from #1~#2 (decimal), for all 0~-1]
+
+Examples:
+# dump.f2fs -i [ino] /dev/sdx
+# dump.f2fs -s 0~-1 /dev/sdx (SIT dump)
+# dump.f2fs -a 0~-1 /dev/sdx (SSA dump)
+
+================================================================================
+DESIGN
+================================================================================
+
+On-disk Layout
+--------------
+
+F2FS divides the whole volume into a number of segments, each of which is fixed
+to 2MB in size. A section is composed of consecutive segments, and a zone
+consists of a set of sections. By default, section and zone sizes are set to one
+segment size identically, but users can easily modify the sizes by mkfs.
+
+F2FS splits the entire volume into six areas, and all the areas except superblock
+consists of multiple segments as described below.
+
+                                            align with the zone size <-|
+                 |-> align with the segment size
+     _________________________________________________________________________
+    |            |            |   Segment   |    Node     |   Segment  |      |
+    | Superblock | Checkpoint |    Info.    |   Address   |   Summary  | Main |
+    |    (SB)    |   (CP)     | Table (SIT) | Table (NAT) | Area (SSA) |      |
+    |____________|_____2______|______N______|______N______|______N_____|__N___|
+                                                                       .      .
+                                                             .                .
+                                                 .                            .
+                                    ._________________________________________.
+                                    |_Segment_|_..._|_Segment_|_..._|_Segment_|
+                                    .           .
+                                    ._________._________
+                                    |_section_|__...__|_
+                                    .            .
+		                    .________.
+	                            |__zone__|
+
+- Superblock (SB)
+ : It is located at the beginning of the partition, and there exist two copies
+   to avoid file system crash. It contains basic partition information and some
+   default parameters of f2fs.
+
+- Checkpoint (CP)
+ : It contains file system information, bitmaps for valid NAT/SIT sets, orphan
+   inode lists, and summary entries of current active segments.
+
+- Segment Information Table (SIT)
+ : It contains segment information such as valid block count and bitmap for the
+   validity of all the blocks.
+
+- Node Address Table (NAT)
+ : It is composed of a block address table for all the node blocks stored in
+   Main area.
+
+- Segment Summary Area (SSA)
+ : It contains summary entries which contains the owner information of all the
+   data and node blocks stored in Main area.
+
+- Main Area
+ : It contains file and directory data including their indices.
+
+In order to avoid misalignment between file system and flash-based storage, F2FS
+aligns the start block address of CP with the segment size. Also, it aligns the
+start block address of Main area with the zone size by reserving some segments
+in SSA area.
+
+Reference the following survey for additional technical details.
+https://wiki.linaro.org/WorkingGroups/Kernel/Projects/FlashCardSurvey
+
+File System Metadata Structure
+------------------------------
+
+F2FS adopts the checkpointing scheme to maintain file system consistency. At
+mount time, F2FS first tries to find the last valid checkpoint data by scanning
+CP area. In order to reduce the scanning time, F2FS uses only two copies of CP.
+One of them always indicates the last valid data, which is called as shadow copy
+mechanism. In addition to CP, NAT and SIT also adopt the shadow copy mechanism.
+
+For file system consistency, each CP points to which NAT and SIT copies are
+valid, as shown as below.
+
+  +--------+----------+---------+
+  |   CP   |    SIT   |   NAT   |
+  +--------+----------+---------+
+  .         .          .          .
+  .            .              .              .
+  .               .                 .                 .
+  +-------+-------+--------+--------+--------+--------+
+  | CP #0 | CP #1 | SIT #0 | SIT #1 | NAT #0 | NAT #1 |
+  +-------+-------+--------+--------+--------+--------+
+     |             ^                          ^
+     |             |                          |
+     `----------------------------------------'
+
+Index Structure
+---------------
+
+The key data structure to manage the data locations is a "node". Similar to
+traditional file structures, F2FS has three types of node: inode, direct node,
+indirect node. F2FS assigns 4KB to an inode block which contains 923 data block
+indices, two direct node pointers, two indirect node pointers, and one double
+indirect node pointer as described below. One direct node block contains 1018
+data blocks, and one indirect node block contains also 1018 node blocks. Thus,
+one inode block (i.e., a file) covers:
+
+  4KB * (923 + 2 * 1018 + 2 * 1018 * 1018 + 1018 * 1018 * 1018) := 3.94TB.
+
+   Inode block (4KB)
+     |- data (923)
+     |- direct node (2)
+     |          `- data (1018)
+     |- indirect node (2)
+     |            `- direct node (1018)
+     |                       `- data (1018)
+     `- double indirect node (1)
+                         `- indirect node (1018)
+			              `- direct node (1018)
+	                                         `- data (1018)
+
+Note that, all the node blocks are mapped by NAT which means the location of
+each node is translated by the NAT table. In the consideration of the wandering
+tree problem, F2FS is able to cut off the propagation of node updates caused by
+leaf data writes.
+
+Directory Structure
+-------------------
+
+A directory entry occupies 11 bytes, which consists of the following attributes.
+
+- hash		hash value of the file name
+- ino		inode number
+- len		the length of file name
+- type		file type such as directory, symlink, etc
+
+A dentry block consists of 214 dentry slots and file names. Therein a bitmap is
+used to represent whether each dentry is valid or not. A dentry block occupies
+4KB with the following composition.
+
+  Dentry Block(4 K) = bitmap (27 bytes) + reserved (3 bytes) +
+	              dentries(11 * 214 bytes) + file name (8 * 214 bytes)
+
+                         [Bucket]
+             +--------------------------------+
+             |dentry block 1 | dentry block 2 |
+             +--------------------------------+
+             .               .
+       .                             .
+  .       [Dentry Block Structure: 4KB]       .
+  +--------+----------+----------+------------+
+  | bitmap | reserved | dentries | file names |
+  +--------+----------+----------+------------+
+  [Dentry Block: 4KB] .   .
+		 .               .
+            .                          .
+            +------+------+-----+------+
+            | hash | ino  | len | type |
+            +------+------+-----+------+
+            [Dentry Structure: 11 bytes]
+
+F2FS implements multi-level hash tables for directory structure. Each level has
+a hash table with dedicated number of hash buckets as shown below. Note that
+"A(2B)" means a bucket includes 2 data blocks.
+
+----------------------
+A : bucket
+B : block
+N : MAX_DIR_HASH_DEPTH
+----------------------
+
+level #0   | A(2B)
+           |
+level #1   | A(2B) - A(2B)
+           |
+level #2   | A(2B) - A(2B) - A(2B) - A(2B)
+     .     |   .       .       .       .
+level #N/2 | A(2B) - A(2B) - A(2B) - A(2B) - A(2B) - ... - A(2B)
+     .     |   .       .       .       .
+level #N   | A(4B) - A(4B) - A(4B) - A(4B) - A(4B) - ... - A(4B)
+
+The number of blocks and buckets are determined by,
+
+                            ,- 2, if n < MAX_DIR_HASH_DEPTH / 2,
+  # of blocks in level #n = |
+                            `- 4, Otherwise
+
+                             ,- 2^(n + dir_level),
+			     |        if n + dir_level < MAX_DIR_HASH_DEPTH / 2,
+  # of buckets in level #n = |
+                             `- 2^((MAX_DIR_HASH_DEPTH / 2) - 1),
+			              Otherwise
+
+When F2FS finds a file name in a directory, at first a hash value of the file
+name is calculated. Then, F2FS scans the hash table in level #0 to find the
+dentry consisting of the file name and its inode number. If not found, F2FS
+scans the next hash table in level #1. In this way, F2FS scans hash tables in
+each levels incrementally from 1 to N. In each levels F2FS needs to scan only
+one bucket determined by the following equation, which shows O(log(# of files))
+complexity.
+
+  bucket number to scan in level #n = (hash value) % (# of buckets in level #n)
+
+In the case of file creation, F2FS finds empty consecutive slots that cover the
+file name. F2FS searches the empty slots in the hash tables of whole levels from
+1 to N in the same way as the lookup operation.
+
+The following figure shows an example of two cases holding children.
+       --------------> Dir <--------------
+       |                                 |
+    child                             child
+
+    child - child                     [hole] - child
+
+    child - child - child             [hole] - [hole] - child
+
+   Case 1:                           Case 2:
+   Number of children = 6,           Number of children = 3,
+   File size = 7                     File size = 7
+
+Default Block Allocation
+------------------------
+
+At runtime, F2FS manages six active logs inside "Main" area: Hot/Warm/Cold node
+and Hot/Warm/Cold data.
+
+- Hot node	contains direct node blocks of directories.
+- Warm node	contains direct node blocks except hot node blocks.
+- Cold node	contains indirect node blocks
+- Hot data	contains dentry blocks
+- Warm data	contains data blocks except hot and cold data blocks
+- Cold data	contains multimedia data or migrated data blocks
+
+LFS has two schemes for free space management: threaded log and copy-and-compac-
+tion. The copy-and-compaction scheme which is known as cleaning, is well-suited
+for devices showing very good sequential write performance, since free segments
+are served all the time for writing new data. However, it suffers from cleaning
+overhead under high utilization. Contrarily, the threaded log scheme suffers
+from random writes, but no cleaning process is needed. F2FS adopts a hybrid
+scheme where the copy-and-compaction scheme is adopted by default, but the
+policy is dynamically changed to the threaded log scheme according to the file
+system status.
+
+In order to align F2FS with underlying flash-based storage, F2FS allocates a
+segment in a unit of section. F2FS expects that the section size would be the
+same as the unit size of garbage collection in FTL. Furthermore, with respect
+to the mapping granularity in FTL, F2FS allocates each section of the active
+logs from different zones as much as possible, since FTL can write the data in
+the active logs into one allocation unit according to its mapping granularity.
+
+Cleaning process
+----------------
+
+F2FS does cleaning both on demand and in the background. On-demand cleaning is
+triggered when there are not enough free segments to serve VFS calls. Background
+cleaner is operated by a kernel thread, and triggers the cleaning job when the
+system is idle.
+
+F2FS supports two victim selection policies: greedy and cost-benefit algorithms.
+In the greedy algorithm, F2FS selects a victim segment having the smallest number
+of valid blocks. In the cost-benefit algorithm, F2FS selects a victim segment
+according to the segment age and the number of valid blocks in order to address
+log block thrashing problem in the greedy algorithm. F2FS adopts the greedy
+algorithm for on-demand cleaner, while background cleaner adopts cost-benefit
+algorithm.
+
+In order to identify whether the data in the victim segment are valid or not,
+F2FS manages a bitmap. Each bit represents the validity of a block, and the
+bitmap is composed of a bit stream covering whole blocks in main area.
+
+Write-hint Policy
+-----------------
+
+1) whint_mode=off. F2FS only passes down WRITE_LIFE_NOT_SET.
+
+2) whint_mode=user-based. F2FS tries to pass down hints given by
+users.
+
+User                  F2FS                     Block
+----                  ----                     -----
+                      META                     WRITE_LIFE_NOT_SET
+                      HOT_NODE                 "
+                      WARM_NODE                "
+                      COLD_NODE                "
+*ioctl(COLD)          COLD_DATA                WRITE_LIFE_EXTREME
+*extension list       "                        "
+
+-- buffered io
+WRITE_LIFE_EXTREME    COLD_DATA                WRITE_LIFE_EXTREME
+WRITE_LIFE_SHORT      HOT_DATA                 WRITE_LIFE_SHORT
+WRITE_LIFE_NOT_SET    WARM_DATA                WRITE_LIFE_NOT_SET
+WRITE_LIFE_NONE       "                        "
+WRITE_LIFE_MEDIUM     "                        "
+WRITE_LIFE_LONG       "                        "
+
+-- direct io
+WRITE_LIFE_EXTREME    COLD_DATA                WRITE_LIFE_EXTREME
+WRITE_LIFE_SHORT      HOT_DATA                 WRITE_LIFE_SHORT
+WRITE_LIFE_NOT_SET    WARM_DATA                WRITE_LIFE_NOT_SET
+WRITE_LIFE_NONE       "                        WRITE_LIFE_NONE
+WRITE_LIFE_MEDIUM     "                        WRITE_LIFE_MEDIUM
+WRITE_LIFE_LONG       "                        WRITE_LIFE_LONG
+
+3) whint_mode=fs-based. F2FS passes down hints with its policy.
+
+User                  F2FS                     Block
+----                  ----                     -----
+                      META                     WRITE_LIFE_MEDIUM;
+                      HOT_NODE                 WRITE_LIFE_NOT_SET
+                      WARM_NODE                "
+                      COLD_NODE                WRITE_LIFE_NONE
+ioctl(COLD)           COLD_DATA                WRITE_LIFE_EXTREME
+extension list        "                        "
+
+-- buffered io
+WRITE_LIFE_EXTREME    COLD_DATA                WRITE_LIFE_EXTREME
+WRITE_LIFE_SHORT      HOT_DATA                 WRITE_LIFE_SHORT
+WRITE_LIFE_NOT_SET    WARM_DATA                WRITE_LIFE_LONG
+WRITE_LIFE_NONE       "                        "
+WRITE_LIFE_MEDIUM     "                        "
+WRITE_LIFE_LONG       "                        "
+
+-- direct io
+WRITE_LIFE_EXTREME    COLD_DATA                WRITE_LIFE_EXTREME
+WRITE_LIFE_SHORT      HOT_DATA                 WRITE_LIFE_SHORT
+WRITE_LIFE_NOT_SET    WARM_DATA                WRITE_LIFE_NOT_SET
+WRITE_LIFE_NONE       "                        WRITE_LIFE_NONE
+WRITE_LIFE_MEDIUM     "                        WRITE_LIFE_MEDIUM
+WRITE_LIFE_LONG       "                        WRITE_LIFE_LONG
diff --git a/Documentation/filesystems/fiemap.txt b/Documentation/filesystems/fiemap.txt
new file mode 100644
index 000000000..f6d9c9910
--- /dev/null
+++ b/Documentation/filesystems/fiemap.txt
@@ -0,0 +1,229 @@
+============
+Fiemap Ioctl
+============
+
+The fiemap ioctl is an efficient method for userspace to get file
+extent mappings. Instead of block-by-block mapping (such as bmap), fiemap
+returns a list of extents.
+
+
+Request Basics
+--------------
+
+A fiemap request is encoded within struct fiemap:
+
+struct fiemap {
+	__u64	fm_start;	 /* logical offset (inclusive) at
+				  * which to start mapping (in) */
+	__u64	fm_length;	 /* logical length of mapping which
+				  * userspace cares about (in) */
+	__u32	fm_flags;	 /* FIEMAP_FLAG_* flags for request (in/out) */
+	__u32	fm_mapped_extents; /* number of extents that were
+				    * mapped (out) */
+	__u32	fm_extent_count; /* size of fm_extents array (in) */
+	__u32	fm_reserved;
+	struct fiemap_extent fm_extents[0]; /* array of mapped extents (out) */
+};
+
+
+fm_start, and fm_length specify the logical range within the file
+which the process would like mappings for. Extents returned mirror
+those on disk - that is, the logical offset of the 1st returned extent
+may start before fm_start, and the range covered by the last returned
+extent may end after fm_length. All offsets and lengths are in bytes.
+
+Certain flags to modify the way in which mappings are looked up can be
+set in fm_flags. If the kernel doesn't understand some particular
+flags, it will return EBADR and the contents of fm_flags will contain
+the set of flags which caused the error. If the kernel is compatible
+with all flags passed, the contents of fm_flags will be unmodified.
+It is up to userspace to determine whether rejection of a particular
+flag is fatal to its operation. This scheme is intended to allow the
+fiemap interface to grow in the future but without losing
+compatibility with old software.
+
+fm_extent_count specifies the number of elements in the fm_extents[] array
+that can be used to return extents.  If fm_extent_count is zero, then the
+fm_extents[] array is ignored (no extents will be returned), and the
+fm_mapped_extents count will hold the number of extents needed in
+fm_extents[] to hold the file's current mapping.  Note that there is
+nothing to prevent the file from changing between calls to FIEMAP.
+
+The following flags can be set in fm_flags:
+
+* FIEMAP_FLAG_SYNC
+If this flag is set, the kernel will sync the file before mapping extents.
+
+* FIEMAP_FLAG_XATTR
+If this flag is set, the extents returned will describe the inodes
+extended attribute lookup tree, instead of its data tree.
+
+
+Extent Mapping
+--------------
+
+Extent information is returned within the embedded fm_extents array
+which userspace must allocate along with the fiemap structure. The
+number of elements in the fiemap_extents[] array should be passed via
+fm_extent_count. The number of extents mapped by kernel will be
+returned via fm_mapped_extents. If the number of fiemap_extents
+allocated is less than would be required to map the requested range,
+the maximum number of extents that can be mapped in the fm_extent[]
+array will be returned and fm_mapped_extents will be equal to
+fm_extent_count. In that case, the last extent in the array will not
+complete the requested range and will not have the FIEMAP_EXTENT_LAST
+flag set (see the next section on extent flags).
+
+Each extent is described by a single fiemap_extent structure as
+returned in fm_extents.
+
+struct fiemap_extent {
+	__u64	fe_logical;  /* logical offset in bytes for the start of
+			      * the extent */
+	__u64	fe_physical; /* physical offset in bytes for the start
+			      * of the extent */
+	__u64	fe_length;   /* length in bytes for the extent */
+	__u64	fe_reserved64[2];
+	__u32	fe_flags;    /* FIEMAP_EXTENT_* flags for this extent */
+	__u32	fe_reserved[3];
+};
+
+All offsets and lengths are in bytes and mirror those on disk.  It is valid
+for an extents logical offset to start before the request or its logical
+length to extend past the request.  Unless FIEMAP_EXTENT_NOT_ALIGNED is
+returned, fe_logical, fe_physical, and fe_length will be aligned to the
+block size of the file system.  With the exception of extents flagged as
+FIEMAP_EXTENT_MERGED, adjacent extents will not be merged.
+
+The fe_flags field contains flags which describe the extent returned.
+A special flag, FIEMAP_EXTENT_LAST is always set on the last extent in
+the file so that the process making fiemap calls can determine when no
+more extents are available, without having to call the ioctl again.
+
+Some flags are intentionally vague and will always be set in the
+presence of other more specific flags. This way a program looking for
+a general property does not have to know all existing and future flags
+which imply that property.
+
+For example, if FIEMAP_EXTENT_DATA_INLINE or FIEMAP_EXTENT_DATA_TAIL
+are set, FIEMAP_EXTENT_NOT_ALIGNED will also be set. A program looking
+for inline or tail-packed data can key on the specific flag. Software
+which simply cares not to try operating on non-aligned extents
+however, can just key on FIEMAP_EXTENT_NOT_ALIGNED, and not have to
+worry about all present and future flags which might imply unaligned
+data. Note that the opposite is not true - it would be valid for
+FIEMAP_EXTENT_NOT_ALIGNED to appear alone.
+
+* FIEMAP_EXTENT_LAST
+This is the last extent in the file. A mapping attempt past this
+extent will return nothing.
+
+* FIEMAP_EXTENT_UNKNOWN
+The location of this extent is currently unknown. This may indicate
+the data is stored on an inaccessible volume or that no storage has
+been allocated for the file yet.
+
+* FIEMAP_EXTENT_DELALLOC
+  - This will also set FIEMAP_EXTENT_UNKNOWN.
+Delayed allocation - while there is data for this extent, its
+physical location has not been allocated yet.
+
+* FIEMAP_EXTENT_ENCODED
+This extent does not consist of plain filesystem blocks but is
+encoded (e.g. encrypted or compressed).  Reading the data in this
+extent via I/O to the block device will have undefined results.
+
+Note that it is *always* undefined to try to update the data
+in-place by writing to the indicated location without the
+assistance of the filesystem, or to access the data using the
+information returned by the FIEMAP interface while the filesystem
+is mounted.  In other words, user applications may only read the
+extent data via I/O to the block device while the filesystem is
+unmounted, and then only if the FIEMAP_EXTENT_ENCODED flag is
+clear; user applications must not try reading or writing to the
+filesystem via the block device under any other circumstances.
+
+* FIEMAP_EXTENT_DATA_ENCRYPTED
+  - This will also set FIEMAP_EXTENT_ENCODED
+The data in this extent has been encrypted by the file system.
+
+* FIEMAP_EXTENT_NOT_ALIGNED
+Extent offsets and length are not guaranteed to be block aligned.
+
+* FIEMAP_EXTENT_DATA_INLINE
+  This will also set FIEMAP_EXTENT_NOT_ALIGNED
+Data is located within a meta data block.
+
+* FIEMAP_EXTENT_DATA_TAIL
+  This will also set FIEMAP_EXTENT_NOT_ALIGNED
+Data is packed into a block with data from other files.
+
+* FIEMAP_EXTENT_UNWRITTEN
+Unwritten extent - the extent is allocated but its data has not been
+initialized.  This indicates the extent's data will be all zero if read
+through the filesystem but the contents are undefined if read directly from
+the device.
+
+* FIEMAP_EXTENT_MERGED
+This will be set when a file does not support extents, i.e., it uses a block
+based addressing scheme.  Since returning an extent for each block back to
+userspace would be highly inefficient, the kernel will try to merge most
+adjacent blocks into 'extents'.
+
+
+VFS -> File System Implementation
+---------------------------------
+
+File systems wishing to support fiemap must implement a ->fiemap callback on
+their inode_operations structure. The fs ->fiemap call is responsible for
+defining its set of supported fiemap flags, and calling a helper function on
+each discovered extent:
+
+struct inode_operations {
+       ...
+
+       int (*fiemap)(struct inode *, struct fiemap_extent_info *, u64 start,
+                     u64 len);
+
+->fiemap is passed struct fiemap_extent_info which describes the
+fiemap request:
+
+struct fiemap_extent_info {
+	unsigned int fi_flags;		/* Flags as passed from user */
+	unsigned int fi_extents_mapped;	/* Number of mapped extents */
+	unsigned int fi_extents_max;	/* Size of fiemap_extent array */
+	struct fiemap_extent *fi_extents_start;	/* Start of fiemap_extent array */
+};
+
+It is intended that the file system should not need to access any of this
+structure directly. Filesystem handlers should be tolerant to signals and return
+EINTR once fatal signal received.
+
+
+Flag checking should be done at the beginning of the ->fiemap callback via the
+fiemap_check_flags() helper:
+
+int fiemap_check_flags(struct fiemap_extent_info *fieinfo, u32 fs_flags);
+
+The struct fieinfo should be passed in as received from ioctl_fiemap(). The
+set of fiemap flags which the fs understands should be passed via fs_flags. If
+fiemap_check_flags finds invalid user flags, it will place the bad values in
+fieinfo->fi_flags and return -EBADR. If the file system gets -EBADR, from
+fiemap_check_flags(), it should immediately exit, returning that error back to
+ioctl_fiemap().
+
+
+For each extent in the request range, the file system should call
+the helper function, fiemap_fill_next_extent():
+
+int fiemap_fill_next_extent(struct fiemap_extent_info *info, u64 logical,
+			    u64 phys, u64 len, u32 flags, u32 dev);
+
+fiemap_fill_next_extent() will use the passed values to populate the
+next free extent in the fm_extents array. 'General' extent flags will
+automatically be set from specific flags on behalf of the calling file
+system so that the userspace API is not broken.
+
+fiemap_fill_next_extent() returns 0 on success, and 1 when the
+user-supplied fm_extents array is full. If an error is encountered
+while copying the extent to user memory, -EFAULT will be returned.
diff --git a/Documentation/filesystems/files.txt b/Documentation/filesystems/files.txt
new file mode 100644
index 000000000..46dfc6b03
--- /dev/null
+++ b/Documentation/filesystems/files.txt
@@ -0,0 +1,123 @@
+File management in the Linux kernel
+-----------------------------------
+
+This document describes how locking for files (struct file)
+and file descriptor table (struct files) works.
+
+Up until 2.6.12, the file descriptor table has been protected
+with a lock (files->file_lock) and reference count (files->count).
+->file_lock protected accesses to all the file related fields
+of the table. ->count was used for sharing the file descriptor
+table between tasks cloned with CLONE_FILES flag. Typically
+this would be the case for posix threads. As with the common
+refcounting model in the kernel, the last task doing
+a put_files_struct() frees the file descriptor (fd) table.
+The files (struct file) themselves are protected using
+reference count (->f_count).
+
+In the new lock-free model of file descriptor management,
+the reference counting is similar, but the locking is
+based on RCU. The file descriptor table contains multiple
+elements - the fd sets (open_fds and close_on_exec, the
+array of file pointers, the sizes of the sets and the array
+etc.). In order for the updates to appear atomic to
+a lock-free reader, all the elements of the file descriptor
+table are in a separate structure - struct fdtable.
+files_struct contains a pointer to struct fdtable through
+which the actual fd table is accessed. Initially the
+fdtable is embedded in files_struct itself. On a subsequent
+expansion of fdtable, a new fdtable structure is allocated
+and files->fdtab points to the new structure. The fdtable
+structure is freed with RCU and lock-free readers either
+see the old fdtable or the new fdtable making the update
+appear atomic. Here are the locking rules for
+the fdtable structure -
+
+1. All references to the fdtable must be done through
+   the files_fdtable() macro :
+
+	struct fdtable *fdt;
+
+	rcu_read_lock();
+
+	fdt = files_fdtable(files);
+	....
+	if (n <= fdt->max_fds)
+		....
+	...
+	rcu_read_unlock();
+
+   files_fdtable() uses rcu_dereference() macro which takes care of
+   the memory barrier requirements for lock-free dereference.
+   The fdtable pointer must be read within the read-side
+   critical section.
+
+2. Reading of the fdtable as described above must be protected
+   by rcu_read_lock()/rcu_read_unlock().
+
+3. For any update to the fd table, files->file_lock must
+   be held.
+
+4. To look up the file structure given an fd, a reader
+   must use either fcheck() or fcheck_files() APIs. These
+   take care of barrier requirements due to lock-free lookup.
+   An example :
+
+	struct file *file;
+
+	rcu_read_lock();
+	file = fcheck(fd);
+	if (file) {
+		...
+	}
+	....
+	rcu_read_unlock();
+
+5. Handling of the file structures is special. Since the look-up
+   of the fd (fget()/fget_light()) are lock-free, it is possible
+   that look-up may race with the last put() operation on the
+   file structure. This is avoided using atomic_long_inc_not_zero()
+   on ->f_count :
+
+	rcu_read_lock();
+	file = fcheck_files(files, fd);
+	if (file) {
+		if (atomic_long_inc_not_zero(&file->f_count))
+			*fput_needed = 1;
+		else
+		/* Didn't get the reference, someone's freed */
+			file = NULL;
+	}
+	rcu_read_unlock();
+	....
+	return file;
+
+   atomic_long_inc_not_zero() detects if refcounts is already zero or
+   goes to zero during increment. If it does, we fail
+   fget()/fget_light().
+
+6. Since both fdtable and file structures can be looked up
+   lock-free, they must be installed using rcu_assign_pointer()
+   API. If they are looked up lock-free, rcu_dereference()
+   must be used. However it is advisable to use files_fdtable()
+   and fcheck()/fcheck_files() which take care of these issues.
+
+7. While updating, the fdtable pointer must be looked up while
+   holding files->file_lock. If ->file_lock is dropped, then
+   another thread expand the files thereby creating a new
+   fdtable and making the earlier fdtable pointer stale.
+   For example :
+
+	spin_lock(&files->file_lock);
+	fd = locate_fd(files, file, start);
+	if (fd >= 0) {
+		/* locate_fd() may have expanded fdtable, load the ptr */
+		fdt = files_fdtable(files);
+		__set_open_fd(fd, fdt);
+		__clear_close_on_exec(fd, fdt);
+		spin_unlock(&files->file_lock);
+	.....
+
+   Since locate_fd() can drop ->file_lock (and reacquire ->file_lock),
+   the fdtable pointer (fdt) must be loaded after locate_fd().
+
diff --git a/Documentation/filesystems/fscrypt.rst b/Documentation/filesystems/fscrypt.rst
new file mode 100644
index 000000000..5b667ee12
--- /dev/null
+++ b/Documentation/filesystems/fscrypt.rst
@@ -0,0 +1,624 @@
+=====================================
+Filesystem-level encryption (fscrypt)
+=====================================
+
+Introduction
+============
+
+fscrypt is a library which filesystems can hook into to support
+transparent encryption of files and directories.
+
+Note: "fscrypt" in this document refers to the kernel-level portion,
+implemented in ``fs/crypto/``, as opposed to the userspace tool
+`fscrypt <https://github.com/google/fscrypt>`_.  This document only
+covers the kernel-level portion.  For command-line examples of how to
+use encryption, see the documentation for the userspace tool `fscrypt
+<https://github.com/google/fscrypt>`_.  Also, it is recommended to use
+the fscrypt userspace tool, or other existing userspace tools such as
+`fscryptctl <https://github.com/google/fscryptctl>`_ or `Android's key
+management system
+<https://source.android.com/security/encryption/file-based>`_, over
+using the kernel's API directly.  Using existing tools reduces the
+chance of introducing your own security bugs.  (Nevertheless, for
+completeness this documentation covers the kernel's API anyway.)
+
+Unlike dm-crypt, fscrypt operates at the filesystem level rather than
+at the block device level.  This allows it to encrypt different files
+with different keys and to have unencrypted files on the same
+filesystem.  This is useful for multi-user systems where each user's
+data-at-rest needs to be cryptographically isolated from the others.
+However, except for filenames, fscrypt does not encrypt filesystem
+metadata.
+
+Unlike eCryptfs, which is a stacked filesystem, fscrypt is integrated
+directly into supported filesystems --- currently ext4, F2FS, and
+UBIFS.  This allows encrypted files to be read and written without
+caching both the decrypted and encrypted pages in the pagecache,
+thereby nearly halving the memory used and bringing it in line with
+unencrypted files.  Similarly, half as many dentries and inodes are
+needed.  eCryptfs also limits encrypted filenames to 143 bytes,
+causing application compatibility issues; fscrypt allows the full 255
+bytes (NAME_MAX).  Finally, unlike eCryptfs, the fscrypt API can be
+used by unprivileged users, with no need to mount anything.
+
+fscrypt does not support encrypting files in-place.  Instead, it
+supports marking an empty directory as encrypted.  Then, after
+userspace provides the key, all regular files, directories, and
+symbolic links created in that directory tree are transparently
+encrypted.
+
+Threat model
+============
+
+Offline attacks
+---------------
+
+Provided that userspace chooses a strong encryption key, fscrypt
+protects the confidentiality of file contents and filenames in the
+event of a single point-in-time permanent offline compromise of the
+block device content.  fscrypt does not protect the confidentiality of
+non-filename metadata, e.g. file sizes, file permissions, file
+timestamps, and extended attributes.  Also, the existence and location
+of holes (unallocated blocks which logically contain all zeroes) in
+files is not protected.
+
+fscrypt is not guaranteed to protect confidentiality or authenticity
+if an attacker is able to manipulate the filesystem offline prior to
+an authorized user later accessing the filesystem.
+
+Online attacks
+--------------
+
+fscrypt (and storage encryption in general) can only provide limited
+protection, if any at all, against online attacks.  In detail:
+
+fscrypt is only resistant to side-channel attacks, such as timing or
+electromagnetic attacks, to the extent that the underlying Linux
+Cryptographic API algorithms are.  If a vulnerable algorithm is used,
+such as a table-based implementation of AES, it may be possible for an
+attacker to mount a side channel attack against the online system.
+Side channel attacks may also be mounted against applications
+consuming decrypted data.
+
+After an encryption key has been provided, fscrypt is not designed to
+hide the plaintext file contents or filenames from other users on the
+same system, regardless of the visibility of the keyring key.
+Instead, existing access control mechanisms such as file mode bits,
+POSIX ACLs, LSMs, or mount namespaces should be used for this purpose.
+Also note that as long as the encryption keys are *anywhere* in
+memory, an online attacker can necessarily compromise them by mounting
+a physical attack or by exploiting any kernel security vulnerability
+which provides an arbitrary memory read primitive.
+
+While it is ostensibly possible to "evict" keys from the system,
+recently accessed encrypted files will remain accessible at least
+until the filesystem is unmounted or the VFS caches are dropped, e.g.
+using ``echo 2 > /proc/sys/vm/drop_caches``.  Even after that, if the
+RAM is compromised before being powered off, it will likely still be
+possible to recover portions of the plaintext file contents, if not
+some of the encryption keys as well.  (Since Linux v4.12, all
+in-kernel keys related to fscrypt are sanitized before being freed.
+However, userspace would need to do its part as well.)
+
+Currently, fscrypt does not prevent a user from maliciously providing
+an incorrect key for another user's existing encrypted files.  A
+protection against this is planned.
+
+Key hierarchy
+=============
+
+Master Keys
+-----------
+
+Each encrypted directory tree is protected by a *master key*.  Master
+keys can be up to 64 bytes long, and must be at least as long as the
+greater of the key length needed by the contents and filenames
+encryption modes being used.  For example, if AES-256-XTS is used for
+contents encryption, the master key must be 64 bytes (512 bits).  Note
+that the XTS mode is defined to require a key twice as long as that
+required by the underlying block cipher.
+
+To "unlock" an encrypted directory tree, userspace must provide the
+appropriate master key.  There can be any number of master keys, each
+of which protects any number of directory trees on any number of
+filesystems.
+
+Userspace should generate master keys either using a cryptographically
+secure random number generator, or by using a KDF (Key Derivation
+Function).  Note that whenever a KDF is used to "stretch" a
+lower-entropy secret such as a passphrase, it is critical that a KDF
+designed for this purpose be used, such as scrypt, PBKDF2, or Argon2.
+
+Per-file keys
+-------------
+
+Master keys are not used to encrypt file contents or names directly.
+Instead, a unique key is derived for each encrypted file, including
+each regular file, directory, and symbolic link.  This has several
+advantages:
+
+- In cryptosystems, the same key material should never be used for
+  different purposes.  Using the master key as both an XTS key for
+  contents encryption and as a CTS-CBC key for filenames encryption
+  would violate this rule.
+- Per-file keys simplify the choice of IVs (Initialization Vectors)
+  for contents encryption.  Without per-file keys, to ensure IV
+  uniqueness both the inode and logical block number would need to be
+  encoded in the IVs.  This would make it impossible to renumber
+  inodes, which e.g. ``resize2fs`` can do when resizing an ext4
+  filesystem.  With per-file keys, it is sufficient to encode just the
+  logical block number in the IVs.
+- Per-file keys strengthen the encryption of filenames, where IVs are
+  reused out of necessity.  With a unique key per directory, IV reuse
+  is limited to within a single directory.
+- Per-file keys allow individual files to be securely erased simply by
+  securely erasing their keys.  (Not yet implemented.)
+
+A KDF (Key Derivation Function) is used to derive per-file keys from
+the master key.  This is done instead of wrapping a randomly-generated
+key for each file because it reduces the size of the encryption xattr,
+which for some filesystems makes the xattr more likely to fit in-line
+in the filesystem's inode table.  With a KDF, only a 16-byte nonce is
+required --- long enough to make key reuse extremely unlikely.  A
+wrapped key, on the other hand, would need to be up to 64 bytes ---
+the length of an AES-256-XTS key.  Furthermore, currently there is no
+requirement to support unlocking a file with multiple alternative
+master keys or to support rotating master keys.  Instead, the master
+keys may be wrapped in userspace, e.g. as done by the `fscrypt
+<https://github.com/google/fscrypt>`_ tool.
+
+The current KDF encrypts the master key using the 16-byte nonce as an
+AES-128-ECB key.  The output is used as the derived key.  If the
+output is longer than needed, then it is truncated to the needed
+length.  Truncation is the norm for directories and symlinks, since
+those use the CTS-CBC encryption mode which requires a key half as
+long as that required by the XTS encryption mode.
+
+Note: this KDF meets the primary security requirement, which is to
+produce unique derived keys that preserve the entropy of the master
+key, assuming that the master key is already a good pseudorandom key.
+However, it is nonstandard and has some problems such as being
+reversible, so it is generally considered to be a mistake!  It may be
+replaced with HKDF or another more standard KDF in the future.
+
+Encryption modes and usage
+==========================
+
+fscrypt allows one encryption mode to be specified for file contents
+and one encryption mode to be specified for filenames.  Different
+directory trees are permitted to use different encryption modes.
+Currently, the following pairs of encryption modes are supported:
+
+- AES-256-XTS for contents and AES-256-CTS-CBC for filenames
+- AES-128-CBC for contents and AES-128-CTS-CBC for filenames
+
+It is strongly recommended to use AES-256-XTS for contents encryption.
+AES-128-CBC was added only for low-powered embedded devices with
+crypto accelerators such as CAAM or CESA that do not support XTS.
+
+New encryption modes can be added relatively easily, without changes
+to individual filesystems.  However, authenticated encryption (AE)
+modes are not currently supported because of the difficulty of dealing
+with ciphertext expansion.
+
+For file contents, each filesystem block is encrypted independently.
+Currently, only the case where the filesystem block size is equal to
+the system's page size (usually 4096 bytes) is supported.  With the
+XTS mode of operation (recommended), the logical block number within
+the file is used as the IV.  With the CBC mode of operation (not
+recommended), ESSIV is used; specifically, the IV for CBC is the
+logical block number encrypted with AES-256, where the AES-256 key is
+the SHA-256 hash of the inode's data encryption key.
+
+For filenames, the full filename is encrypted at once.  Because of the
+requirements to retain support for efficient directory lookups and
+filenames of up to 255 bytes, a constant initialization vector (IV) is
+used.  However, each encrypted directory uses a unique key, which
+limits IV reuse to within a single directory.  Note that IV reuse in
+the context of CTS-CBC encryption means that when the original
+filenames share a common prefix at least as long as the cipher block
+size (16 bytes for AES), the corresponding encrypted filenames will
+also share a common prefix.  This is undesirable; it may be fixed in
+the future by switching to an encryption mode that is a strong
+pseudorandom permutation on arbitrary-length messages, e.g. the HEH
+(Hash-Encrypt-Hash) mode.
+
+Since filenames are encrypted with the CTS-CBC mode of operation, the
+plaintext and ciphertext filenames need not be multiples of the AES
+block size, i.e. 16 bytes.  However, the minimum size that can be
+encrypted is 16 bytes, so shorter filenames are NUL-padded to 16 bytes
+before being encrypted.  In addition, to reduce leakage of filename
+lengths via their ciphertexts, all filenames are NUL-padded to the
+next 4, 8, 16, or 32-byte boundary (configurable).  32 is recommended
+since this provides the best confidentiality, at the cost of making
+directory entries consume slightly more space.  Note that since NUL
+(``\0``) is not otherwise a valid character in filenames, the padding
+will never produce duplicate plaintexts.
+
+Symbolic link targets are considered a type of filename and are
+encrypted in the same way as filenames in directory entries.  Each
+symlink also uses a unique key; hence, the hardcoded IV is not a
+problem for symlinks.
+
+User API
+========
+
+Setting an encryption policy
+----------------------------
+
+The FS_IOC_SET_ENCRYPTION_POLICY ioctl sets an encryption policy on an
+empty directory or verifies that a directory or regular file already
+has the specified encryption policy.  It takes in a pointer to a
+:c:type:`struct fscrypt_policy`, defined as follows::
+
+    #define FS_KEY_DESCRIPTOR_SIZE  8
+
+    struct fscrypt_policy {
+            __u8 version;
+            __u8 contents_encryption_mode;
+            __u8 filenames_encryption_mode;
+            __u8 flags;
+            __u8 master_key_descriptor[FS_KEY_DESCRIPTOR_SIZE];
+    };
+
+This structure must be initialized as follows:
+
+- ``version`` must be 0.
+
+- ``contents_encryption_mode`` and ``filenames_encryption_mode`` must
+  be set to constants from ``<linux/fs.h>`` which identify the
+  encryption modes to use.  If unsure, use
+  FS_ENCRYPTION_MODE_AES_256_XTS (1) for ``contents_encryption_mode``
+  and FS_ENCRYPTION_MODE_AES_256_CTS (4) for
+  ``filenames_encryption_mode``.
+
+- ``flags`` must be set to a value from ``<linux/fs.h>`` which
+  identifies the amount of NUL-padding to use when encrypting
+  filenames.  If unsure, use FS_POLICY_FLAGS_PAD_32 (0x3).
+
+- ``master_key_descriptor`` specifies how to find the master key in
+  the keyring; see `Adding keys`_.  It is up to userspace to choose a
+  unique ``master_key_descriptor`` for each master key.  The e4crypt
+  and fscrypt tools use the first 8 bytes of
+  ``SHA-512(SHA-512(master_key))``, but this particular scheme is not
+  required.  Also, the master key need not be in the keyring yet when
+  FS_IOC_SET_ENCRYPTION_POLICY is executed.  However, it must be added
+  before any files can be created in the encrypted directory.
+
+If the file is not yet encrypted, then FS_IOC_SET_ENCRYPTION_POLICY
+verifies that the file is an empty directory.  If so, the specified
+encryption policy is assigned to the directory, turning it into an
+encrypted directory.  After that, and after providing the
+corresponding master key as described in `Adding keys`_, all regular
+files, directories (recursively), and symlinks created in the
+directory will be encrypted, inheriting the same encryption policy.
+The filenames in the directory's entries will be encrypted as well.
+
+Alternatively, if the file is already encrypted, then
+FS_IOC_SET_ENCRYPTION_POLICY validates that the specified encryption
+policy exactly matches the actual one.  If they match, then the ioctl
+returns 0.  Otherwise, it fails with EEXIST.  This works on both
+regular files and directories, including nonempty directories.
+
+Note that the ext4 filesystem does not allow the root directory to be
+encrypted, even if it is empty.  Users who want to encrypt an entire
+filesystem with one key should consider using dm-crypt instead.
+
+FS_IOC_SET_ENCRYPTION_POLICY can fail with the following errors:
+
+- ``EACCES``: the file is not owned by the process's uid, nor does the
+  process have the CAP_FOWNER capability in a namespace with the file
+  owner's uid mapped
+- ``EEXIST``: the file is already encrypted with an encryption policy
+  different from the one specified
+- ``EINVAL``: an invalid encryption policy was specified (invalid
+  version, mode(s), or flags)
+- ``ENOTDIR``: the file is unencrypted and is a regular file, not a
+  directory
+- ``ENOTEMPTY``: the file is unencrypted and is a nonempty directory
+- ``ENOTTY``: this type of filesystem does not implement encryption
+- ``EOPNOTSUPP``: the kernel was not configured with encryption
+  support for this filesystem, or the filesystem superblock has not
+  had encryption enabled on it.  (For example, to use encryption on an
+  ext4 filesystem, CONFIG_EXT4_ENCRYPTION must be enabled in the
+  kernel config, and the superblock must have had the "encrypt"
+  feature flag enabled using ``tune2fs -O encrypt`` or ``mkfs.ext4 -O
+  encrypt``.)
+- ``EPERM``: this directory may not be encrypted, e.g. because it is
+  the root directory of an ext4 filesystem
+- ``EROFS``: the filesystem is readonly
+
+Getting an encryption policy
+----------------------------
+
+The FS_IOC_GET_ENCRYPTION_POLICY ioctl retrieves the :c:type:`struct
+fscrypt_policy`, if any, for a directory or regular file.  See above
+for the struct definition.  No additional permissions are required
+beyond the ability to open the file.
+
+FS_IOC_GET_ENCRYPTION_POLICY can fail with the following errors:
+
+- ``EINVAL``: the file is encrypted, but it uses an unrecognized
+  encryption context format
+- ``ENODATA``: the file is not encrypted
+- ``ENOTTY``: this type of filesystem does not implement encryption
+- ``EOPNOTSUPP``: the kernel was not configured with encryption
+  support for this filesystem
+
+Note: if you only need to know whether a file is encrypted or not, on
+most filesystems it is also possible to use the FS_IOC_GETFLAGS ioctl
+and check for FS_ENCRYPT_FL, or to use the statx() system call and
+check for STATX_ATTR_ENCRYPTED in stx_attributes.
+
+Getting the per-filesystem salt
+-------------------------------
+
+Some filesystems, such as ext4 and F2FS, also support the deprecated
+ioctl FS_IOC_GET_ENCRYPTION_PWSALT.  This ioctl retrieves a randomly
+generated 16-byte value stored in the filesystem superblock.  This
+value is intended to used as a salt when deriving an encryption key
+from a passphrase or other low-entropy user credential.
+
+FS_IOC_GET_ENCRYPTION_PWSALT is deprecated.  Instead, prefer to
+generate and manage any needed salt(s) in userspace.
+
+Adding keys
+-----------
+
+To provide a master key, userspace must add it to an appropriate
+keyring using the add_key() system call (see:
+``Documentation/security/keys/core.rst``).  The key type must be
+"logon"; keys of this type are kept in kernel memory and cannot be
+read back by userspace.  The key description must be "fscrypt:"
+followed by the 16-character lower case hex representation of the
+``master_key_descriptor`` that was set in the encryption policy.  The
+key payload must conform to the following structure::
+
+    #define FS_MAX_KEY_SIZE 64
+
+    struct fscrypt_key {
+            u32 mode;
+            u8 raw[FS_MAX_KEY_SIZE];
+            u32 size;
+    };
+
+``mode`` is ignored; just set it to 0.  The actual key is provided in
+``raw`` with ``size`` indicating its size in bytes.  That is, the
+bytes ``raw[0..size-1]`` (inclusive) are the actual key.
+
+The key description prefix "fscrypt:" may alternatively be replaced
+with a filesystem-specific prefix such as "ext4:".  However, the
+filesystem-specific prefixes are deprecated and should not be used in
+new programs.
+
+There are several different types of keyrings in which encryption keys
+may be placed, such as a session keyring, a user session keyring, or a
+user keyring.  Each key must be placed in a keyring that is "attached"
+to all processes that might need to access files encrypted with it, in
+the sense that request_key() will find the key.  Generally, if only
+processes belonging to a specific user need to access a given
+encrypted directory and no session keyring has been installed, then
+that directory's key should be placed in that user's user session
+keyring or user keyring.  Otherwise, a session keyring should be
+installed if needed, and the key should be linked into that session
+keyring, or in a keyring linked into that session keyring.
+
+Note: introducing the complex visibility semantics of keyrings here
+was arguably a mistake --- especially given that by design, after any
+process successfully opens an encrypted file (thereby setting up the
+per-file key), possessing the keyring key is not actually required for
+any process to read/write the file until its in-memory inode is
+evicted.  In the future there probably should be a way to provide keys
+directly to the filesystem instead, which would make the intended
+semantics clearer.
+
+Access semantics
+================
+
+With the key
+------------
+
+With the encryption key, encrypted regular files, directories, and
+symlinks behave very similarly to their unencrypted counterparts ---
+after all, the encryption is intended to be transparent.  However,
+astute users may notice some differences in behavior:
+
+- Unencrypted files, or files encrypted with a different encryption
+  policy (i.e. different key, modes, or flags), cannot be renamed or
+  linked into an encrypted directory; see `Encryption policy
+  enforcement`_.  Attempts to do so will fail with EXDEV.  However,
+  encrypted files can be renamed within an encrypted directory, or
+  into an unencrypted directory.
+
+  Note: "moving" an unencrypted file into an encrypted directory, e.g.
+  with the `mv` program, is implemented in userspace by a copy
+  followed by a delete.  Be aware that the original unencrypted data
+  may remain recoverable from free space on the disk; prefer to keep
+  all files encrypted from the very beginning.  The `shred` program
+  may be used to overwrite the source files but isn't guaranteed to be
+  effective on all filesystems and storage devices.
+
+- Direct I/O is not supported on encrypted files.  Attempts to use
+  direct I/O on such files will fall back to buffered I/O.
+
+- The fallocate operations FALLOC_FL_COLLAPSE_RANGE,
+  FALLOC_FL_INSERT_RANGE, and FALLOC_FL_ZERO_RANGE are not supported
+  on encrypted files and will fail with EOPNOTSUPP.
+
+- Online defragmentation of encrypted files is not supported.  The
+  EXT4_IOC_MOVE_EXT and F2FS_IOC_MOVE_RANGE ioctls will fail with
+  EOPNOTSUPP.
+
+- The ext4 filesystem does not support data journaling with encrypted
+  regular files.  It will fall back to ordered data mode instead.
+
+- DAX (Direct Access) is not supported on encrypted files.
+
+- The st_size of an encrypted symlink will not necessarily give the
+  length of the symlink target as required by POSIX.  It will actually
+  give the length of the ciphertext, which will be slightly longer
+  than the plaintext due to NUL-padding and an extra 2-byte overhead.
+
+- The maximum length of an encrypted symlink is 2 bytes shorter than
+  the maximum length of an unencrypted symlink.  For example, on an
+  EXT4 filesystem with a 4K block size, unencrypted symlinks can be up
+  to 4095 bytes long, while encrypted symlinks can only be up to 4093
+  bytes long (both lengths excluding the terminating null).
+
+Note that mmap *is* supported.  This is possible because the pagecache
+for an encrypted file contains the plaintext, not the ciphertext.
+
+Without the key
+---------------
+
+Some filesystem operations may be performed on encrypted regular
+files, directories, and symlinks even before their encryption key has
+been provided:
+
+- File metadata may be read, e.g. using stat().
+
+- Directories may be listed, in which case the filenames will be
+  listed in an encoded form derived from their ciphertext.  The
+  current encoding algorithm is described in `Filename hashing and
+  encoding`_.  The algorithm is subject to change, but it is
+  guaranteed that the presented filenames will be no longer than
+  NAME_MAX bytes, will not contain the ``/`` or ``\0`` characters, and
+  will uniquely identify directory entries.
+
+  The ``.`` and ``..`` directory entries are special.  They are always
+  present and are not encrypted or encoded.
+
+- Files may be deleted.  That is, nondirectory files may be deleted
+  with unlink() as usual, and empty directories may be deleted with
+  rmdir() as usual.  Therefore, ``rm`` and ``rm -r`` will work as
+  expected.
+
+- Symlink targets may be read and followed, but they will be presented
+  in encrypted form, similar to filenames in directories.  Hence, they
+  are unlikely to point to anywhere useful.
+
+Without the key, regular files cannot be opened or truncated.
+Attempts to do so will fail with ENOKEY.  This implies that any
+regular file operations that require a file descriptor, such as
+read(), write(), mmap(), fallocate(), and ioctl(), are also forbidden.
+
+Also without the key, files of any type (including directories) cannot
+be created or linked into an encrypted directory, nor can a name in an
+encrypted directory be the source or target of a rename, nor can an
+O_TMPFILE temporary file be created in an encrypted directory.  All
+such operations will fail with ENOKEY.
+
+It is not currently possible to backup and restore encrypted files
+without the encryption key.  This would require special APIs which
+have not yet been implemented.
+
+Encryption policy enforcement
+=============================
+
+After an encryption policy has been set on a directory, all regular
+files, directories, and symbolic links created in that directory
+(recursively) will inherit that encryption policy.  Special files ---
+that is, named pipes, device nodes, and UNIX domain sockets --- will
+not be encrypted.
+
+Except for those special files, it is forbidden to have unencrypted
+files, or files encrypted with a different encryption policy, in an
+encrypted directory tree.  Attempts to link or rename such a file into
+an encrypted directory will fail with EXDEV.  This is also enforced
+during ->lookup() to provide limited protection against offline
+attacks that try to disable or downgrade encryption in known locations
+where applications may later write sensitive data.  It is recommended
+that systems implementing a form of "verified boot" take advantage of
+this by validating all top-level encryption policies prior to access.
+
+Implementation details
+======================
+
+Encryption context
+------------------
+
+An encryption policy is represented on-disk by a :c:type:`struct
+fscrypt_context`.  It is up to individual filesystems to decide where
+to store it, but normally it would be stored in a hidden extended
+attribute.  It should *not* be exposed by the xattr-related system
+calls such as getxattr() and setxattr() because of the special
+semantics of the encryption xattr.  (In particular, there would be
+much confusion if an encryption policy were to be added to or removed
+from anything other than an empty directory.)  The struct is defined
+as follows::
+
+    #define FS_KEY_DESCRIPTOR_SIZE  8
+    #define FS_KEY_DERIVATION_NONCE_SIZE 16
+
+    struct fscrypt_context {
+            u8 format;
+            u8 contents_encryption_mode;
+            u8 filenames_encryption_mode;
+            u8 flags;
+            u8 master_key_descriptor[FS_KEY_DESCRIPTOR_SIZE];
+            u8 nonce[FS_KEY_DERIVATION_NONCE_SIZE];
+    };
+
+Note that :c:type:`struct fscrypt_context` contains the same
+information as :c:type:`struct fscrypt_policy` (see `Setting an
+encryption policy`_), except that :c:type:`struct fscrypt_context`
+also contains a nonce.  The nonce is randomly generated by the kernel
+and is used to derive the inode's encryption key as described in
+`Per-file keys`_.
+
+Data path changes
+-----------------
+
+For the read path (->readpage()) of regular files, filesystems can
+read the ciphertext into the page cache and decrypt it in-place.  The
+page lock must be held until decryption has finished, to prevent the
+page from becoming visible to userspace prematurely.
+
+For the write path (->writepage()) of regular files, filesystems
+cannot encrypt data in-place in the page cache, since the cached
+plaintext must be preserved.  Instead, filesystems must encrypt into a
+temporary buffer or "bounce page", then write out the temporary
+buffer.  Some filesystems, such as UBIFS, already use temporary
+buffers regardless of encryption.  Other filesystems, such as ext4 and
+F2FS, have to allocate bounce pages specially for encryption.
+
+Filename hashing and encoding
+-----------------------------
+
+Modern filesystems accelerate directory lookups by using indexed
+directories.  An indexed directory is organized as a tree keyed by
+filename hashes.  When a ->lookup() is requested, the filesystem
+normally hashes the filename being looked up so that it can quickly
+find the corresponding directory entry, if any.
+
+With encryption, lookups must be supported and efficient both with and
+without the encryption key.  Clearly, it would not work to hash the
+plaintext filenames, since the plaintext filenames are unavailable
+without the key.  (Hashing the plaintext filenames would also make it
+impossible for the filesystem's fsck tool to optimize encrypted
+directories.)  Instead, filesystems hash the ciphertext filenames,
+i.e. the bytes actually stored on-disk in the directory entries.  When
+asked to do a ->lookup() with the key, the filesystem just encrypts
+the user-supplied name to get the ciphertext.
+
+Lookups without the key are more complicated.  The raw ciphertext may
+contain the ``\0`` and ``/`` characters, which are illegal in
+filenames.  Therefore, readdir() must base64-encode the ciphertext for
+presentation.  For most filenames, this works fine; on ->lookup(), the
+filesystem just base64-decodes the user-supplied name to get back to
+the raw ciphertext.
+
+However, for very long filenames, base64 encoding would cause the
+filename length to exceed NAME_MAX.  To prevent this, readdir()
+actually presents long filenames in an abbreviated form which encodes
+a strong "hash" of the ciphertext filename, along with the optional
+filesystem-specific hash(es) needed for directory lookups.  This
+allows the filesystem to still, with a high degree of confidence, map
+the filename given in ->lookup() back to a particular directory entry
+that was previously listed by readdir().  See :c:type:`struct
+fscrypt_digested_name` in the source for more details.
+
+Note that the precise way that filenames are presented to userspace
+without the key is subject to change in the future.  It is only meant
+as a way to temporarily present valid filenames so that commands like
+``rm -r`` work as expected on encrypted directories.
diff --git a/Documentation/filesystems/fuse-io.txt b/Documentation/filesystems/fuse-io.txt
new file mode 100644
index 000000000..07b8f73f1
--- /dev/null
+++ b/Documentation/filesystems/fuse-io.txt
@@ -0,0 +1,38 @@
+Fuse supports the following I/O modes:
+
+- direct-io
+- cached
+  + write-through
+  + writeback-cache
+
+The direct-io mode can be selected with the FOPEN_DIRECT_IO flag in the
+FUSE_OPEN reply.
+
+In direct-io mode the page cache is completely bypassed for reads and writes.
+No read-ahead takes place. Shared mmap is disabled.
+
+In cached mode reads may be satisfied from the page cache, and data may be
+read-ahead by the kernel to fill the cache.  The cache is always kept consistent
+after any writes to the file.  All mmap modes are supported.
+
+The cached mode has two sub modes controlling how writes are handled.  The
+write-through mode is the default and is supported on all kernels.  The
+writeback-cache mode may be selected by the FUSE_WRITEBACK_CACHE flag in the
+FUSE_INIT reply.
+
+In write-through mode each write is immediately sent to userspace as one or more
+WRITE requests, as well as updating any cached pages (and caching previously
+uncached, but fully written pages).  No READ requests are ever sent for writes,
+so when an uncached page is partially written, the page is discarded.
+
+In writeback-cache mode (enabled by the FUSE_WRITEBACK_CACHE flag) writes go to
+the cache only, which means that the write(2) syscall can often complete very
+fast.  Dirty pages are written back implicitly (background writeback or page
+reclaim on memory pressure) or explicitly (invoked by close(2), fsync(2) and
+when the last ref to the file is being released on munmap(2)).  This mode
+assumes that all changes to the filesystem go through the FUSE kernel module
+(size and atime/ctime/mtime attributes are kept up-to-date by the kernel), so
+it's generally not suitable for network filesystems.  If a partial page is
+written, then the page needs to be first read from userspace.  This means, that
+even for files opened for O_WRONLY it is possible that READ requests will be
+generated by the kernel.
diff --git a/Documentation/filesystems/fuse.txt b/Documentation/filesystems/fuse.txt
new file mode 100644
index 000000000..13af4a49e
--- /dev/null
+++ b/Documentation/filesystems/fuse.txt
@@ -0,0 +1,423 @@
+Definitions
+~~~~~~~~~~~
+
+Userspace filesystem:
+
+  A filesystem in which data and metadata are provided by an ordinary
+  userspace process.  The filesystem can be accessed normally through
+  the kernel interface.
+
+Filesystem daemon:
+
+  The process(es) providing the data and metadata of the filesystem.
+
+Non-privileged mount (or user mount):
+
+  A userspace filesystem mounted by a non-privileged (non-root) user.
+  The filesystem daemon is running with the privileges of the mounting
+  user.  NOTE: this is not the same as mounts allowed with the "user"
+  option in /etc/fstab, which is not discussed here.
+
+Filesystem connection:
+
+  A connection between the filesystem daemon and the kernel.  The
+  connection exists until either the daemon dies, or the filesystem is
+  umounted.  Note that detaching (or lazy umounting) the filesystem
+  does _not_ break the connection, in this case it will exist until
+  the last reference to the filesystem is released.
+
+Mount owner:
+
+  The user who does the mounting.
+
+User:
+
+  The user who is performing filesystem operations.
+
+What is FUSE?
+~~~~~~~~~~~~~
+
+FUSE is a userspace filesystem framework.  It consists of a kernel
+module (fuse.ko), a userspace library (libfuse.*) and a mount utility
+(fusermount).
+
+One of the most important features of FUSE is allowing secure,
+non-privileged mounts.  This opens up new possibilities for the use of
+filesystems.  A good example is sshfs: a secure network filesystem
+using the sftp protocol.
+
+The userspace library and utilities are available from the FUSE
+homepage:
+
+  http://fuse.sourceforge.net/
+
+Filesystem type
+~~~~~~~~~~~~~~~
+
+The filesystem type given to mount(2) can be one of the following:
+
+'fuse'
+
+  This is the usual way to mount a FUSE filesystem.  The first
+  argument of the mount system call may contain an arbitrary string,
+  which is not interpreted by the kernel.
+
+'fuseblk'
+
+  The filesystem is block device based.  The first argument of the
+  mount system call is interpreted as the name of the device.
+
+Mount options
+~~~~~~~~~~~~~
+
+'fd=N'
+
+  The file descriptor to use for communication between the userspace
+  filesystem and the kernel.  The file descriptor must have been
+  obtained by opening the FUSE device ('/dev/fuse').
+
+'rootmode=M'
+
+  The file mode of the filesystem's root in octal representation.
+
+'user_id=N'
+
+  The numeric user id of the mount owner.
+
+'group_id=N'
+
+  The numeric group id of the mount owner.
+
+'default_permissions'
+
+  By default FUSE doesn't check file access permissions, the
+  filesystem is free to implement its access policy or leave it to
+  the underlying file access mechanism (e.g. in case of network
+  filesystems).  This option enables permission checking, restricting
+  access based on file mode.  It is usually useful together with the
+  'allow_other' mount option.
+
+'allow_other'
+
+  This option overrides the security measure restricting file access
+  to the user mounting the filesystem.  This option is by default only
+  allowed to root, but this restriction can be removed with a
+  (userspace) configuration option.
+
+'max_read=N'
+
+  With this option the maximum size of read operations can be set.
+  The default is infinite.  Note that the size of read requests is
+  limited anyway to 32 pages (which is 128kbyte on i386).
+
+'blksize=N'
+
+  Set the block size for the filesystem.  The default is 512.  This
+  option is only valid for 'fuseblk' type mounts.
+
+Control filesystem
+~~~~~~~~~~~~~~~~~~
+
+There's a control filesystem for FUSE, which can be mounted by:
+
+  mount -t fusectl none /sys/fs/fuse/connections
+
+Mounting it under the '/sys/fs/fuse/connections' directory makes it
+backwards compatible with earlier versions.
+
+Under the fuse control filesystem each connection has a directory
+named by a unique number.
+
+For each connection the following files exist within this directory:
+
+ 'waiting'
+
+  The number of requests which are waiting to be transferred to
+  userspace or being processed by the filesystem daemon.  If there is
+  no filesystem activity and 'waiting' is non-zero, then the
+  filesystem is hung or deadlocked.
+
+ 'abort'
+
+  Writing anything into this file will abort the filesystem
+  connection.  This means that all waiting requests will be aborted an
+  error returned for all aborted and new requests.
+
+Only the owner of the mount may read or write these files.
+
+Interrupting filesystem operations
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+If a process issuing a FUSE filesystem request is interrupted, the
+following will happen:
+
+  1) If the request is not yet sent to userspace AND the signal is
+     fatal (SIGKILL or unhandled fatal signal), then the request is
+     dequeued and returns immediately.
+
+  2) If the request is not yet sent to userspace AND the signal is not
+     fatal, then an 'interrupted' flag is set for the request.  When
+     the request has been successfully transferred to userspace and
+     this flag is set, an INTERRUPT request is queued.
+
+  3) If the request is already sent to userspace, then an INTERRUPT
+     request is queued.
+
+INTERRUPT requests take precedence over other requests, so the
+userspace filesystem will receive queued INTERRUPTs before any others.
+
+The userspace filesystem may ignore the INTERRUPT requests entirely,
+or may honor them by sending a reply to the _original_ request, with
+the error set to EINTR.
+
+It is also possible that there's a race between processing the
+original request and its INTERRUPT request.  There are two possibilities:
+
+  1) The INTERRUPT request is processed before the original request is
+     processed
+
+  2) The INTERRUPT request is processed after the original request has
+     been answered
+
+If the filesystem cannot find the original request, it should wait for
+some timeout and/or a number of new requests to arrive, after which it
+should reply to the INTERRUPT request with an EAGAIN error.  In case
+1) the INTERRUPT request will be requeued.  In case 2) the INTERRUPT
+reply will be ignored.
+
+Aborting a filesystem connection
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+It is possible to get into certain situations where the filesystem is
+not responding.  Reasons for this may be:
+
+  a) Broken userspace filesystem implementation
+
+  b) Network connection down
+
+  c) Accidental deadlock
+
+  d) Malicious deadlock
+
+(For more on c) and d) see later sections)
+
+In either of these cases it may be useful to abort the connection to
+the filesystem.  There are several ways to do this:
+
+  - Kill the filesystem daemon.  Works in case of a) and b)
+
+  - Kill the filesystem daemon and all users of the filesystem.  Works
+    in all cases except some malicious deadlocks
+
+  - Use forced umount (umount -f).  Works in all cases but only if
+    filesystem is still attached (it hasn't been lazy unmounted)
+
+  - Abort filesystem through the FUSE control filesystem.  Most
+    powerful method, always works.
+
+How do non-privileged mounts work?
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+Since the mount() system call is a privileged operation, a helper
+program (fusermount) is needed, which is installed setuid root.
+
+The implication of providing non-privileged mounts is that the mount
+owner must not be able to use this capability to compromise the
+system.  Obvious requirements arising from this are:
+
+ A) mount owner should not be able to get elevated privileges with the
+    help of the mounted filesystem
+
+ B) mount owner should not get illegitimate access to information from
+    other users' and the super user's processes
+
+ C) mount owner should not be able to induce undesired behavior in
+    other users' or the super user's processes
+
+How are requirements fulfilled?
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+ A) The mount owner could gain elevated privileges by either:
+
+     1) creating a filesystem containing a device file, then opening
+	this device
+
+     2) creating a filesystem containing a suid or sgid application,
+	then executing this application
+
+    The solution is not to allow opening device files and ignore
+    setuid and setgid bits when executing programs.  To ensure this
+    fusermount always adds "nosuid" and "nodev" to the mount options
+    for non-privileged mounts.
+
+ B) If another user is accessing files or directories in the
+    filesystem, the filesystem daemon serving requests can record the
+    exact sequence and timing of operations performed.  This
+    information is otherwise inaccessible to the mount owner, so this
+    counts as an information leak.
+
+    The solution to this problem will be presented in point 2) of C).
+
+ C) There are several ways in which the mount owner can induce
+    undesired behavior in other users' processes, such as:
+
+     1) mounting a filesystem over a file or directory which the mount
+        owner could otherwise not be able to modify (or could only
+        make limited modifications).
+
+        This is solved in fusermount, by checking the access
+        permissions on the mountpoint and only allowing the mount if
+        the mount owner can do unlimited modification (has write
+        access to the mountpoint, and mountpoint is not a "sticky"
+        directory)
+
+     2) Even if 1) is solved the mount owner can change the behavior
+        of other users' processes.
+
+         i) It can slow down or indefinitely delay the execution of a
+           filesystem operation creating a DoS against the user or the
+           whole system.  For example a suid application locking a
+           system file, and then accessing a file on the mount owner's
+           filesystem could be stopped, and thus causing the system
+           file to be locked forever.
+
+         ii) It can present files or directories of unlimited length, or
+           directory structures of unlimited depth, possibly causing a
+           system process to eat up diskspace, memory or other
+           resources, again causing DoS.
+
+	The solution to this as well as B) is not to allow processes
+	to access the filesystem, which could otherwise not be
+	monitored or manipulated by the mount owner.  Since if the
+	mount owner can ptrace a process, it can do all of the above
+	without using a FUSE mount, the same criteria as used in
+	ptrace can be used to check if a process is allowed to access
+	the filesystem or not.
+
+	Note that the ptrace check is not strictly necessary to
+	prevent B/2/i, it is enough to check if mount owner has enough
+	privilege to send signal to the process accessing the
+	filesystem, since SIGSTOP can be used to get a similar effect.
+
+I think these limitations are unacceptable?
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+If a sysadmin trusts the users enough, or can ensure through other
+measures, that system processes will never enter non-privileged
+mounts, it can relax the last limitation with a "user_allow_other"
+config option.  If this config option is set, the mounting user can
+add the "allow_other" mount option which disables the check for other
+users' processes.
+
+Kernel - userspace interface
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+The following diagram shows how a filesystem operation (in this
+example unlink) is performed in FUSE.
+
+NOTE: everything in this description is greatly simplified
+
+ |  "rm /mnt/fuse/file"               |  FUSE filesystem daemon
+ |                                    |
+ |                                    |  >sys_read()
+ |                                    |    >fuse_dev_read()
+ |                                    |      >request_wait()
+ |                                    |        [sleep on fc->waitq]
+ |                                    |
+ |  >sys_unlink()                     |
+ |    >fuse_unlink()                  |
+ |      [get request from             |
+ |       fc->unused_list]             |
+ |      >request_send()               |
+ |        [queue req on fc->pending]  |
+ |        [wake up fc->waitq]         |        [woken up]
+ |        >request_wait_answer()      |
+ |          [sleep on req->waitq]     |
+ |                                    |      <request_wait()
+ |                                    |      [remove req from fc->pending]
+ |                                    |      [copy req to read buffer]
+ |                                    |      [add req to fc->processing]
+ |                                    |    <fuse_dev_read()
+ |                                    |  <sys_read()
+ |                                    |
+ |                                    |  [perform unlink]
+ |                                    |
+ |                                    |  >sys_write()
+ |                                    |    >fuse_dev_write()
+ |                                    |      [look up req in fc->processing]
+ |                                    |      [remove from fc->processing]
+ |                                    |      [copy write buffer to req]
+ |          [woken up]                |      [wake up req->waitq]
+ |                                    |    <fuse_dev_write()
+ |                                    |  <sys_write()
+ |        <request_wait_answer()      |
+ |      <request_send()               |
+ |      [add request to               |
+ |       fc->unused_list]             |
+ |    <fuse_unlink()                  |
+ |  <sys_unlink()                     |
+
+There are a couple of ways in which to deadlock a FUSE filesystem.
+Since we are talking about unprivileged userspace programs,
+something must be done about these.
+
+Scenario 1 -  Simple deadlock
+-----------------------------
+
+ |  "rm /mnt/fuse/file"               |  FUSE filesystem daemon
+ |                                    |
+ |  >sys_unlink("/mnt/fuse/file")     |
+ |    [acquire inode semaphore        |
+ |     for "file"]                    |
+ |    >fuse_unlink()                  |
+ |      [sleep on req->waitq]         |
+ |                                    |  <sys_read()
+ |                                    |  >sys_unlink("/mnt/fuse/file")
+ |                                    |    [acquire inode semaphore
+ |                                    |     for "file"]
+ |                                    |    *DEADLOCK*
+
+The solution for this is to allow the filesystem to be aborted.
+
+Scenario 2 - Tricky deadlock
+----------------------------
+
+This one needs a carefully crafted filesystem.  It's a variation on
+the above, only the call back to the filesystem is not explicit,
+but is caused by a pagefault.
+
+ |  Kamikaze filesystem thread 1      |  Kamikaze filesystem thread 2
+ |                                    |
+ |  [fd = open("/mnt/fuse/file")]     |  [request served normally]
+ |  [mmap fd to 'addr']               |
+ |  [close fd]                        |  [FLUSH triggers 'magic' flag]
+ |  [read a byte from addr]           |
+ |    >do_page_fault()                |
+ |      [find or create page]         |
+ |      [lock page]                   |
+ |      >fuse_readpage()              |
+ |         [queue READ request]       |
+ |         [sleep on req->waitq]      |
+ |                                    |  [read request to buffer]
+ |                                    |  [create reply header before addr]
+ |                                    |  >sys_write(addr - headerlength)
+ |                                    |    >fuse_dev_write()
+ |                                    |      [look up req in fc->processing]
+ |                                    |      [remove from fc->processing]
+ |                                    |      [copy write buffer to req]
+ |                                    |        >do_page_fault()
+ |                                    |           [find or create page]
+ |                                    |           [lock page]
+ |                                    |           * DEADLOCK *
+
+Solution is basically the same as above.
+
+An additional problem is that while the write buffer is being copied
+to the request, the request must not be interrupted/aborted.  This is
+because the destination address of the copy may not be valid after the
+request has returned.
+
+This is solved with doing the copy atomically, and allowing abort
+while the page(s) belonging to the write buffer are faulted with
+get_user_pages().  The 'req->locked' flag indicates when the copy is
+taking place, and abort is delayed until this flag is unset.
diff --git a/Documentation/filesystems/gfs2-glocks.txt b/Documentation/filesystems/gfs2-glocks.txt
new file mode 100644
index 000000000..705962363
--- /dev/null
+++ b/Documentation/filesystems/gfs2-glocks.txt
@@ -0,0 +1,232 @@
+                   Glock internal locking rules
+                  ------------------------------
+
+This documents the basic principles of the glock state machine
+internals. Each glock (struct gfs2_glock in fs/gfs2/incore.h)
+has two main (internal) locks:
+
+ 1. A spinlock (gl_lockref.lock) which protects the internal state such
+    as gl_state, gl_target and the list of holders (gl_holders)
+ 2. A non-blocking bit lock, GLF_LOCK, which is used to prevent other
+    threads from making calls to the DLM, etc. at the same time. If a
+    thread takes this lock, it must then call run_queue (usually via the
+    workqueue) when it releases it in order to ensure any pending tasks
+    are completed.
+
+The gl_holders list contains all the queued lock requests (not
+just the holders) associated with the glock. If there are any
+held locks, then they will be contiguous entries at the head
+of the list. Locks are granted in strictly the order that they
+are queued, except for those marked LM_FLAG_PRIORITY which are
+used only during recovery, and even then only for journal locks.
+
+There are three lock states that users of the glock layer can request,
+namely shared (SH), deferred (DF) and exclusive (EX). Those translate
+to the following DLM lock modes:
+
+Glock mode    | DLM lock mode
+------------------------------
+    UN        |    IV/NL  Unlocked (no DLM lock associated with glock) or NL
+    SH        |    PR     (Protected read)
+    DF        |    CW     (Concurrent write)
+    EX        |    EX     (Exclusive)
+
+Thus DF is basically a shared mode which is incompatible with the "normal"
+shared lock mode, SH. In GFS2 the DF mode is used exclusively for direct I/O
+operations. The glocks are basically a lock plus some routines which deal
+with cache management. The following rules apply for the cache:
+
+Glock mode   |  Cache data | Cache Metadata | Dirty Data | Dirty Metadata
+--------------------------------------------------------------------------
+    UN       |     No      |       No       |     No     |      No
+    SH       |     Yes     |       Yes      |     No     |      No
+    DF       |     No      |       Yes      |     No     |      No
+    EX       |     Yes     |       Yes      |     Yes    |      Yes
+
+These rules are implemented using the various glock operations which
+are defined for each type of glock. Not all types of glocks use
+all the modes. Only inode glocks use the DF mode for example.
+
+Table of glock operations and per type constants:
+
+Field            | Purpose
+----------------------------------------------------------------------------
+go_xmote_th      | Called before remote state change (e.g. to sync dirty data)
+go_xmote_bh      | Called after remote state change (e.g. to refill cache)
+go_inval         | Called if remote state change requires invalidating the cache
+go_demote_ok     | Returns boolean value of whether its ok to demote a glock
+                 | (e.g. checks timeout, and that there is no cached data)
+go_lock          | Called for the first local holder of a lock
+go_unlock        | Called on the final local unlock of a lock
+go_dump          | Called to print content of object for debugfs file, or on
+                 | error to dump glock to the log.
+go_type          | The type of the glock, LM_TYPE_.....
+go_callback	 | Called if the DLM sends a callback to drop this lock
+go_flags	 | GLOF_ASPACE is set, if the glock has an address space
+                 | associated with it
+
+The minimum hold time for each lock is the time after a remote lock
+grant for which we ignore remote demote requests. This is in order to
+prevent a situation where locks are being bounced around the cluster
+from node to node with none of the nodes making any progress. This
+tends to show up most with shared mmaped files which are being written
+to by multiple nodes. By delaying the demotion in response to a
+remote callback, that gives the userspace program time to make
+some progress before the pages are unmapped.
+
+There is a plan to try and remove the go_lock and go_unlock callbacks
+if possible, in order to try and speed up the fast path though the locking.
+Also, eventually we hope to make the glock "EX" mode locally shared
+such that any local locking will be done with the i_mutex as required
+rather than via the glock.
+
+Locking rules for glock operations:
+
+Operation     |  GLF_LOCK bit lock held |  gl_lockref.lock spinlock held
+-------------------------------------------------------------------------
+go_xmote_th   |       Yes               |       No
+go_xmote_bh   |       Yes               |       No
+go_inval      |       Yes               |       No
+go_demote_ok  |       Sometimes         |       Yes
+go_lock       |       Yes               |       No
+go_unlock     |       Yes               |       No
+go_dump       |       Sometimes         |       Yes
+go_callback   |       Sometimes (N/A)   |       Yes
+
+N.B. Operations must not drop either the bit lock or the spinlock
+if its held on entry. go_dump and do_demote_ok must never block.
+Note that go_dump will only be called if the glock's state
+indicates that it is caching uptodate data.
+
+Glock locking order within GFS2:
+
+ 1. i_rwsem (if required)
+ 2. Rename glock (for rename only)
+ 3. Inode glock(s)
+    (Parents before children, inodes at "same level" with same parent in
+     lock number order)
+ 4. Rgrp glock(s) (for (de)allocation operations)
+ 5. Transaction glock (via gfs2_trans_begin) for non-read operations
+ 6. i_rw_mutex (if required)
+ 7. Page lock  (always last, very important!)
+
+There are two glocks per inode. One deals with access to the inode
+itself (locking order as above), and the other, known as the iopen
+glock is used in conjunction with the i_nlink field in the inode to
+determine the lifetime of the inode in question. Locking of inodes
+is on a per-inode basis. Locking of rgrps is on a per rgrp basis.
+In general we prefer to lock local locks prior to cluster locks.
+
+                            Glock Statistics
+                           ------------------
+
+The stats are divided into two sets: those relating to the
+super block and those relating to an individual glock. The
+super block stats are done on a per cpu basis in order to
+try and reduce the overhead of gathering them. They are also
+further divided by glock type. All timings are in nanoseconds.
+
+In the case of both the super block and glock statistics,
+the same information is gathered in each case. The super
+block timing statistics are used to provide default values for
+the glock timing statistics, so that newly created glocks
+should have, as far as possible, a sensible starting point.
+The per-glock counters are initialised to zero when the
+glock is created. The per-glock statistics are lost when
+the glock is ejected from memory.
+
+The statistics are divided into three pairs of mean and
+variance, plus two counters. The mean/variance pairs are
+smoothed exponential estimates and the algorithm used is
+one which will be very familiar to those used to calculation
+of round trip times in network code. See "TCP/IP Illustrated,
+Volume 1", W. Richard Stevens, sect 21.3, "Round-Trip Time Measurement",
+p. 299 and onwards. Also, Volume 2, Sect. 25.10, p. 838 and onwards.
+Unlike the TCP/IP Illustrated case, the mean and variance are
+not scaled, but are in units of integer nanoseconds.
+
+The three pairs of mean/variance measure the following
+things:
+
+ 1. DLM lock time (non-blocking requests)
+ 2. DLM lock time (blocking requests)
+ 3. Inter-request time (again to the DLM)
+
+A non-blocking request is one which will complete right
+away, whatever the state of the DLM lock in question. That
+currently means any requests when (a) the current state of
+the lock is exclusive, i.e. a lock demotion (b) the requested
+state is either null or unlocked (again, a demotion) or (c) the
+"try lock" flag is set. A blocking request covers all the other
+lock requests.
+
+There are two counters. The first is there primarily to show
+how many lock requests have been made, and thus how much data
+has gone into the mean/variance calculations. The other counter
+is counting queuing of holders at the top layer of the glock
+code. Hopefully that number will be a lot larger than the number
+of dlm lock requests issued.
+
+So why gather these statistics? There are several reasons
+we'd like to get a better idea of these timings:
+
+1. To be able to better set the glock "min hold time"
+2. To spot performance issues more easily
+3. To improve the algorithm for selecting resource groups for
+allocation (to base it on lock wait time, rather than blindly
+using a "try lock")
+
+Due to the smoothing action of the updates, a step change in
+some input quantity being sampled will only fully be taken
+into account after 8 samples (or 4 for the variance) and this
+needs to be carefully considered when interpreting the
+results.
+
+Knowing both the time it takes a lock request to complete and
+the average time between lock requests for a glock means we
+can compute the total percentage of the time for which the
+node is able to use a glock vs. time that the rest of the
+cluster has its share. That will be very useful when setting
+the lock min hold time.
+
+Great care has been taken to ensure that we
+measure exactly the quantities that we want, as accurately
+as possible. There are always inaccuracies in any
+measuring system, but I hope this is as accurate as we
+can reasonably make it.
+
+Per sb stats can be found here:
+/sys/kernel/debug/gfs2/<fsname>/sbstats
+Per glock stats can be found here:
+/sys/kernel/debug/gfs2/<fsname>/glstats
+
+Assuming that debugfs is mounted on /sys/kernel/debug and also
+that <fsname> is replaced with the name of the gfs2 filesystem
+in question.
+
+The abbreviations used in the output as are follows:
+
+srtt     - Smoothed round trip time for non-blocking dlm requests
+srttvar  - Variance estimate for srtt
+srttb    - Smoothed round trip time for (potentially) blocking dlm requests
+srttvarb - Variance estimate for srttb
+sirt     - Smoothed inter-request time (for dlm requests)
+sirtvar  - Variance estimate for sirt
+dlm      - Number of dlm requests made (dcnt in glstats file)
+queue    - Number of glock requests queued (qcnt in glstats file)
+
+The sbstats file contains a set of these stats for each glock type (so 8 lines
+for each type) and for each cpu (one column per cpu). The glstats file contains
+a set of these stats for each glock in a similar format to the glocks file, but
+using the format mean/variance for each of the timing stats.
+
+The gfs2_glock_lock_time tracepoint prints out the current values of the stats
+for the glock in question, along with some addition information on each dlm
+reply that is received:
+
+status - The status of the dlm request
+flags  - The dlm request flags
+tdiff  - The time taken by this specific request
+(remaining fields as per above list)
+
+
diff --git a/Documentation/filesystems/gfs2-uevents.txt b/Documentation/filesystems/gfs2-uevents.txt
new file mode 100644
index 000000000..19a19ebeb
--- /dev/null
+++ b/Documentation/filesystems/gfs2-uevents.txt
@@ -0,0 +1,100 @@
+                              uevents and GFS2
+                             ==================
+
+During the lifetime of a GFS2 mount, a number of uevents are generated.
+This document explains what the events are and what they are used
+for (by gfs_controld in gfs2-utils).
+
+A list of GFS2 uevents
+-----------------------
+
+1. ADD
+
+The ADD event occurs at mount time. It will always be the first
+uevent generated by the newly created filesystem. If the mount
+is successful, an ONLINE uevent will follow.  If it is not successful
+then a REMOVE uevent will follow.
+
+The ADD uevent has two environment variables: SPECTATOR=[0|1]
+and RDONLY=[0|1] that specify the spectator status (a read-only mount
+with no journal assigned), and read-only (with journal assigned) status
+of the filesystem respectively.
+
+2. ONLINE
+
+The ONLINE uevent is generated after a successful mount or remount. It
+has the same environment variables as the ADD uevent. The ONLINE
+uevent, along with the two environment variables for spectator and
+RDONLY are a relatively recent addition (2.6.32-rc+) and will not
+be generated by older kernels.
+
+3. CHANGE
+
+The CHANGE uevent is used in two places. One is when reporting the
+successful mount of the filesystem by the first node (FIRSTMOUNT=Done).
+This is used as a signal by gfs_controld that it is then ok for other
+nodes in the cluster to mount the filesystem.
+
+The other CHANGE uevent is used to inform of the completion
+of journal recovery for one of the filesystems journals. It has
+two environment variables, JID= which specifies the journal id which
+has just been recovered, and RECOVERY=[Done|Failed] to indicate the
+success (or otherwise) of the operation. These uevents are generated
+for every journal recovered, whether it is during the initial mount
+process or as the result of gfs_controld requesting a specific journal
+recovery via the /sys/fs/gfs2/<fsname>/lock_module/recovery file.
+
+Because the CHANGE uevent was used (in early versions of gfs_controld)
+without checking the environment variables to discover the state, we
+cannot add any more functions to it without running the risk of
+someone using an older version of the user tools and breaking their
+cluster. For this reason the ONLINE uevent was used when adding a new
+uevent for a successful mount or remount.
+
+4. OFFLINE
+
+The OFFLINE uevent is only generated due to filesystem errors and is used
+as part of the "withdraw" mechanism. Currently this doesn't give any
+information about what the error is, which is something that needs to
+be fixed.
+
+5. REMOVE
+
+The REMOVE uevent is generated at the end of an unsuccessful mount
+or at the end of a umount of the filesystem. All REMOVE uevents will
+have been preceded by at least an ADD uevent for the same filesystem,
+and unlike the other uevents is generated automatically by the kernel's
+kobject subsystem.
+
+
+Information common to all GFS2 uevents (uevent environment variables)
+----------------------------------------------------------------------
+
+1. LOCKTABLE=
+
+The LOCKTABLE is a string, as supplied on the mount command
+line (locktable=) or via fstab. It is used as a filesystem label
+as well as providing the information for a lock_dlm mount to be
+able to join the cluster.
+
+2. LOCKPROTO=
+
+The LOCKPROTO is a string, and its value depends on what is set
+on the mount command line, or via fstab. It will be either
+lock_nolock or lock_dlm. In the future other lock managers
+may be supported.
+
+3. JOURNALID=
+
+If a journal is in use by the filesystem (journals are not
+assigned for spectator mounts) then this will give the
+numeric journal id in all GFS2 uevents.
+
+4. UUID=
+
+With recent versions of gfs2-utils, mkfs.gfs2 writes a UUID
+into the filesystem superblock. If it exists, this will
+be included in every uevent relating to the filesystem.
+
+
+
diff --git a/Documentation/filesystems/gfs2.txt b/Documentation/filesystems/gfs2.txt
new file mode 100644
index 000000000..cc4f23066
--- /dev/null
+++ b/Documentation/filesystems/gfs2.txt
@@ -0,0 +1,45 @@
+Global File System
+------------------
+
+https://fedorahosted.org/cluster/wiki/HomePage
+
+GFS is a cluster file system. It allows a cluster of computers to
+simultaneously use a block device that is shared between them (with FC,
+iSCSI, NBD, etc).  GFS reads and writes to the block device like a local
+file system, but also uses a lock module to allow the computers coordinate
+their I/O so file system consistency is maintained.  One of the nifty
+features of GFS is perfect consistency -- changes made to the file system
+on one machine show up immediately on all other machines in the cluster.
+
+GFS uses interchangeable inter-node locking mechanisms, the currently
+supported mechanisms are:
+
+  lock_nolock -- allows gfs to be used as a local file system
+
+  lock_dlm -- uses a distributed lock manager (dlm) for inter-node locking
+  The dlm is found at linux/fs/dlm/
+
+Lock_dlm depends on user space cluster management systems found
+at the URL above.
+
+To use gfs as a local file system, no external clustering systems are
+needed, simply:
+
+  $ mkfs -t gfs2 -p lock_nolock -j 1 /dev/block_device
+  $ mount -t gfs2 /dev/block_device /dir
+
+If you are using Fedora, you need to install the gfs2-utils package
+and, for lock_dlm, you will also need to install the cman package
+and write a cluster.conf as per the documentation. For F17 and above
+cman has been replaced by the dlm package.
+
+GFS2 is not on-disk compatible with previous versions of GFS, but it
+is pretty close.
+
+The following man pages can be found at the URL above:
+  fsck.gfs2		to repair a filesystem
+  gfs2_grow		to expand a filesystem online
+  gfs2_jadd		to add journals to a filesystem online
+  tunegfs2		to manipulate, examine and tune a filesystem
+  gfs2_convert	to convert a gfs filesystem to gfs2 in-place
+  mkfs.gfs2		to make a filesystem
diff --git a/Documentation/filesystems/hfs.txt b/Documentation/filesystems/hfs.txt
new file mode 100644
index 000000000..d096df6db
--- /dev/null
+++ b/Documentation/filesystems/hfs.txt
@@ -0,0 +1,82 @@
+Note: This filesystem doesn't have a maintainer.
+
+Macintosh HFS Filesystem for Linux
+==================================
+
+HFS stands for ``Hierarchical File System'' and is the filesystem used
+by the Mac Plus and all later Macintosh models.  Earlier Macintosh
+models used MFS (``Macintosh File System''), which is not supported,
+MacOS 8.1 and newer support a filesystem called HFS+ that's similar to
+HFS but is extended in various areas.  Use the hfsplus filesystem driver
+to access such filesystems from Linux.
+
+
+Mount options
+=============
+
+When mounting an HFS filesystem, the following options are accepted:
+
+  creator=cccc, type=cccc
+	Specifies the creator/type values as shown by the MacOS finder
+	used for creating new files.  Default values: '????'.
+
+  uid=n, gid=n
+  	Specifies the user/group that owns all files on the filesystems.
+	Default:  user/group id of the mounting process.
+
+  dir_umask=n, file_umask=n, umask=n
+	Specifies the umask used for all files , all directories or all
+	files and directories.  Defaults to the umask of the mounting process.
+
+  session=n
+  	Select the CDROM session to mount as HFS filesystem.  Defaults to
+	leaving that decision to the CDROM driver.  This option will fail
+	with anything but a CDROM as underlying devices.
+
+  part=n
+  	Select partition number n from the devices.  Does only makes
+	sense for CDROMS because they can't be partitioned under Linux.
+	For disk devices the generic partition parsing code does this
+	for us.  Defaults to not parsing the partition table at all.
+
+  quiet
+  	Ignore invalid mount options instead of complaining.
+
+
+Writing to HFS Filesystems
+==========================
+
+HFS is not a UNIX filesystem, thus it does not have the usual features you'd
+expect:
+
+ o You can't modify the set-uid, set-gid, sticky or executable bits or the uid
+   and gid of files.
+ o You can't create hard- or symlinks, device files, sockets or FIFOs.
+
+HFS does on the other have the concepts of multiple forks per file.  These
+non-standard forks are represented as hidden additional files in the normal
+filesystems namespace which is kind of a cludge and makes the semantics for
+the a little strange:
+
+ o You can't create, delete or rename resource forks of files or the
+   Finder's metadata.
+ o They are however created (with default values), deleted and renamed
+   along with the corresponding data fork or directory.
+ o Copying files to a different filesystem will loose those attributes
+   that are essential for MacOS to work.
+
+
+Creating HFS filesystems
+===================================
+
+The hfsutils package from Robert Leslie contains a program called
+hformat that can be used to create HFS filesystem. See
+<http://www.mars.org/home/rob/proj/hfs/> for details.
+
+
+Credits
+=======
+
+The HFS drivers was written by Paul H. Hargrovea (hargrove@sccm.Stanford.EDU).
+Roman Zippel (roman@ardistech.com) rewrote large parts of the code and brought
+in btree routines derived from Brad Boyer's hfsplus driver.
diff --git a/Documentation/filesystems/hfsplus.txt b/Documentation/filesystems/hfsplus.txt
new file mode 100644
index 000000000..59f7569fc
--- /dev/null
+++ b/Documentation/filesystems/hfsplus.txt
@@ -0,0 +1,59 @@
+
+Macintosh HFSPlus Filesystem for Linux
+======================================
+
+HFSPlus is a filesystem first introduced in MacOS 8.1.
+HFSPlus has several extensions to HFS, including 32-bit allocation
+blocks, 255-character unicode filenames, and file sizes of 2^63 bytes.
+
+
+Mount options
+=============
+
+When mounting an HFSPlus filesystem, the following options are accepted:
+
+  creator=cccc, type=cccc
+	Specifies the creator/type values as shown by the MacOS finder
+	used for creating new files.  Default values: '????'.
+
+  uid=n, gid=n
+	Specifies the user/group that owns all files on the filesystem
+	that have uninitialized permissions structures.
+	Default:  user/group id of the mounting process.
+
+  umask=n
+	Specifies the umask (in octal) used for files and directories
+	that have uninitialized permissions structures.
+	Default:  umask of the mounting process.
+
+  session=n
+	Select the CDROM session to mount as HFSPlus filesystem.  Defaults to
+	leaving that decision to the CDROM driver.  This option will fail
+	with anything but a CDROM as underlying devices.
+
+  part=n
+	Select partition number n from the devices.  This option only makes
+	sense for CDROMs because they can't be partitioned under Linux.
+	For disk devices the generic partition parsing code does this
+	for us.  Defaults to not parsing the partition table at all.
+
+  decompose
+	Decompose file name characters.
+
+  nodecompose
+	Do not decompose file name characters.
+
+  force
+	Used to force write access to volumes that are marked as journalled
+	or locked.  Use at your own risk.
+
+  nls=cccc
+	Encoding to use when presenting file names.
+
+
+References
+==========
+
+kernel source:		<file:fs/hfsplus>
+
+Apple Technote 1150	https://developer.apple.com/legacy/library/technotes/tn/tn1150.html
diff --git a/Documentation/filesystems/hpfs.txt b/Documentation/filesystems/hpfs.txt
new file mode 100644
index 000000000..74630bd50
--- /dev/null
+++ b/Documentation/filesystems/hpfs.txt
@@ -0,0 +1,296 @@
+Read/Write HPFS 2.09
+1998-2004, Mikulas Patocka
+
+email: mikulas@artax.karlin.mff.cuni.cz
+homepage: http://artax.karlin.mff.cuni.cz/~mikulas/vyplody/hpfs/index-e.cgi
+
+CREDITS:
+Chris Smith, 1993, original read-only HPFS, some code and hpfs structures file
+	is taken from it
+Jacques Gelinas, MSDos mmap, Inspired by fs/nfs/mmap.c (Jon Tombs 15 Aug 1993)
+Werner Almesberger, 1992, 1993, MSDos option parser & CR/LF conversion
+
+Mount options
+
+uid=xxx,gid=xxx,umask=xxx (default uid=gid=0 umask=default_system_umask)
+	Set owner/group/mode for files that do not have it specified in extended
+	attributes. Mode is inverted umask - for example umask 027 gives owner
+	all permission, group read permission and anybody else no access. Note
+	that for files mode is anded with 0666. If you want files to have 'x'
+	rights, you must use extended attributes.
+case=lower,asis (default asis)
+	File name lowercasing in readdir.
+conv=binary,text,auto (default binary)
+	CR/LF -> LF conversion, if auto, decision is made according to extension
+	- there is a list of text extensions (I thing it's better to not convert
+	text file than to damage binary file). If you want to change that list,
+	change it in the source. Original readonly HPFS contained some strange
+	heuristic algorithm that I removed. I thing it's danger to let the
+	computer decide whether file is text or binary. For example, DJGPP
+	binaries contain small text message at the beginning and they could be
+	misidentified and damaged under some circumstances.
+check=none,normal,strict (default normal)
+	Check level. Selecting none will cause only little speedup and big
+	danger. I tried to write it so that it won't crash if check=normal on
+	corrupted filesystems. check=strict means many superfluous checks -
+	used for debugging (for example it checks if file is allocated in
+	bitmaps when accessing it).
+errors=continue,remount-ro,panic (default remount-ro)
+	Behaviour when filesystem errors found.
+chkdsk=no,errors,always (default errors)
+	When to mark filesystem dirty so that OS/2 checks it.
+eas=no,ro,rw (default rw)
+	What to do with extended attributes. 'no' - ignore them and use always
+	values specified in uid/gid/mode options. 'ro' - read extended
+	attributes but do not create them. 'rw' - create extended attributes
+	when you use chmod/chown/chgrp/mknod/ln -s on the filesystem.
+timeshift=(-)nnn (default 0)
+	Shifts the time by nnn seconds. For example, if you see under linux
+	one hour more, than under os/2, use timeshift=-3600.
+
+
+File names
+
+As in OS/2, filenames are case insensitive. However, shell thinks that names
+are case sensitive, so for example when you create a file FOO, you can use
+'cat FOO', 'cat Foo', 'cat foo' or 'cat F*' but not 'cat f*'. Note, that you
+also won't be able to compile linux kernel (and maybe other things) on HPFS
+because kernel creates different files with names like bootsect.S and
+bootsect.s. When searching for file thats name has characters >= 128, codepages
+are used - see below.
+OS/2 ignores dots and spaces at the end of file name, so this driver does as
+well. If you create 'a. ...', the file 'a' will be created, but you can still
+access it under names 'a.', 'a..', 'a .  . . ' etc.
+
+
+Extended attributes
+
+On HPFS partitions, OS/2 can associate to each file a special information called
+extended attributes. Extended attributes are pairs of (key,value) where key is
+an ascii string identifying that attribute and value is any string of bytes of
+variable length. OS/2 stores window and icon positions and file types there. So
+why not use it for unix-specific info like file owner or access rights? This
+driver can do it. If you chown/chgrp/chmod on a hpfs partition, extended
+attributes with keys "UID", "GID" or "MODE" and 2-byte values are created. Only
+that extended attributes those value differs from defaults specified in mount
+options are created. Once created, the extended attributes are never deleted,
+they're just changed. It means that when your default uid=0 and you type
+something like 'chown luser file; chown root file' the file will contain
+extended attribute UID=0. And when you umount the fs and mount it again with
+uid=luser_uid, the file will be still owned by root! If you chmod file to 444,
+extended attribute "MODE" will not be set, this special case is done by setting
+read-only flag. When you mknod a block or char device, besides "MODE", the
+special 4-byte extended attribute "DEV" will be created containing the device
+number. Currently this driver cannot resize extended attributes - it means
+that if somebody (I don't know who?) has set "UID", "GID", "MODE" or "DEV"
+attributes with different sizes, they won't be rewritten and changing these
+values doesn't work.
+
+
+Symlinks
+
+You can do symlinks on HPFS partition, symlinks are achieved by setting extended
+attribute named "SYMLINK" with symlink value. Like on ext2, you can chown and
+chgrp symlinks but I don't know what is it good for. chmoding symlink results
+in chmoding file where symlink points. These symlinks are just for Linux use and
+incompatible with OS/2. OS/2 PmShell symlinks are not supported because they are
+stored in very crazy way. They tried to do it so that link changes when file is
+moved ... sometimes it works. But the link is partly stored in directory
+extended attributes and partly in OS2SYS.INI. I don't want (and don't know how)
+to analyze or change OS2SYS.INI.
+
+
+Codepages
+
+HPFS can contain several uppercasing tables for several codepages and each
+file has a pointer to codepage its name is in. However OS/2 was created in
+America where people don't care much about codepages and so multiple codepages
+support is quite buggy. I have Czech OS/2 working in codepage 852 on my disk.
+Once I booted English OS/2 working in cp 850 and I created a file on my 852
+partition. It marked file name codepage as 850 - good. But when I again booted
+Czech OS/2, the file was completely inaccessible under any name. It seems that
+OS/2 uppercases the search pattern with its system code page (852) and file
+name it's comparing to with its code page (850). These could never match. Is it
+really what IBM developers wanted? But problems continued. When I created in
+Czech OS/2 another file in that directory, that file was inaccessible too. OS/2
+probably uses different uppercasing method when searching where to place a file
+(note, that files in HPFS directory must be sorted) and when searching for
+a file. Finally when I opened this directory in PmShell, PmShell crashed (the
+funny thing was that, when rebooted, PmShell tried to reopen this directory
+again :-). chkdsk happily ignores these errors and only low-level disk
+modification saved me.  Never mix different language versions of OS/2 on one
+system although HPFS was designed to allow that.
+OK, I could implement complex codepage support to this driver but I think it
+would cause more problems than benefit with such buggy implementation in OS/2.
+So this driver simply uses first codepage it finds for uppercasing and
+lowercasing no matter what's file codepage index. Usually all file names are in
+this codepage - if you don't try to do what I described above :-)
+
+
+Known bugs
+
+HPFS386 on OS/2 server is not supported. HPFS386 installed on normal OS/2 client
+should work. If you have OS/2 server, use only read-only mode. I don't know how
+to handle some HPFS386 structures like access control list or extended perm
+list, I don't know how to delete them when file is deleted and how to not
+overwrite them with extended attributes. Send me some info on these structures
+and I'll make it. However, this driver should detect presence of HPFS386
+structures, remount read-only and not destroy them (I hope).
+
+When there's not enough space for extended attributes, they will be truncated
+and no error is returned.
+
+OS/2 can't access files if the path is longer than about 256 chars but this
+driver allows you to do it. chkdsk ignores such errors.
+
+Sometimes you won't be able to delete some files on a very full filesystem
+(returning error ENOSPC). That's because file in non-leaf node in directory tree
+(one directory, if it's large, has dirents in tree on HPFS) must be replaced
+with another node when deleted. And that new file might have larger name than
+the old one so the new name doesn't fit in directory node (dnode). And that
+would result in directory tree splitting, that takes disk space. Workaround is
+to delete other files that are leaf (probability that the file is non-leaf is
+about 1/50) or to truncate file first to make some space.
+You encounter this problem only if you have many directories so that
+preallocated directory band is full i.e.
+	number_of_directories / size_of_filesystem_in_mb > 4.
+
+You can't delete open directories.
+
+You can't rename over directories (what is it good for?).
+
+Renaming files so that only case changes doesn't work. This driver supports it
+but vfs doesn't. Something like 'mv file FILE' won't work.
+
+All atimes and directory mtimes are not updated. That's because of performance
+reasons. If you extremely wish to update them, let me know, I'll write it (but
+it will be slow).
+
+When the system is out of memory and swap, it may slightly corrupt filesystem
+(lost files, unbalanced directories). (I guess all filesystem may do it).
+
+When compiled, you get warning: function declaration isn't a prototype. Does
+anybody know what does it mean?
+
+
+What does "unbalanced tree" message mean?
+
+Old versions of this driver created sometimes unbalanced dnode trees. OS/2
+chkdsk doesn't scream if the tree is unbalanced (and sometimes creates
+unbalanced trees too :-) but both HPFS and HPFS386 contain bug that it rarely
+crashes when the tree is not balanced. This driver handles unbalanced trees
+correctly and writes warning if it finds them. If you see this message, this is
+probably because of directories created with old version of this driver.
+Workaround is to move all files from that directory to another and then back
+again. Do it in Linux, not OS/2! If you see this message in directory that is
+whole created by this driver, it is BUG - let me know about it.
+
+
+Bugs in OS/2
+
+When you have two (or more) lost directories pointing each to other, chkdsk
+locks up when repairing filesystem.
+
+Sometimes (I think it's random) when you create a file with one-char name under
+OS/2, OS/2 marks it as 'long'. chkdsk then removes this flag saying "Minor fs
+error corrected".
+
+File names like "a .b" are marked as 'long' by OS/2 but chkdsk "corrects" it and
+marks them as short (and writes "minor fs error corrected"). This bug is not in
+HPFS386.
+
+Codepage bugs described above.
+
+If you don't install fixpacks, there are many, many more...
+
+
+History
+
+0.90 First public release
+0.91 Fixed bug that caused shooting to memory when write_inode was called on
+	open inode (rarely happened)
+0.92 Fixed a little memory leak in freeing directory inodes
+0.93 Fixed bug that locked up the machine when there were too many filenames
+	with first 15 characters same
+     Fixed write_file to zero file when writing behind file end
+0.94 Fixed a little memory leak when trying to delete busy file or directory
+0.95 Fixed a bug that i_hpfs_parent_dir was not updated when moving files
+1.90 First version for 2.1.1xx kernels
+1.91 Fixed a bug that chk_sectors failed when sectors were at the end of disk
+     Fixed a race-condition when write_inode is called while deleting file
+     Fixed a bug that could possibly happen (with very low probability) when
+     	using 0xff in filenames
+     Rewritten locking to avoid race-conditions
+     Mount option 'eas' now works
+     Fsync no longer returns error
+     Files beginning with '.' are marked hidden
+     Remount support added
+     Alloc is not so slow when filesystem becomes full
+     Atimes are no more updated because it slows down operation
+     Code cleanup (removed all commented debug prints)
+1.92 Corrected a bug when sync was called just before closing file
+1.93 Modified, so that it works with kernels >= 2.1.131, I don't know if it
+	works with previous versions
+     Fixed a possible problem with disks > 64G (but I don't have one, so I can't
+     	test it)
+     Fixed a file overflow at 2G
+     Added new option 'timeshift'
+     Changed behaviour on HPFS386: It is now possible to operate on HPFS386 in
+     	read-only mode
+     Fixed a bug that slowed down alloc and prevented allocating 100% space
+     	(this bug was not destructive)
+1.94 Added workaround for one bug in Linux
+     Fixed one buffer leak
+     Fixed some incompatibilities with large extended attributes (but it's still
+	not 100% ok, I have no info on it and OS/2 doesn't want to create them)
+     Rewritten allocation
+     Fixed a bug with i_blocks (du sometimes didn't display correct values)
+     Directories have no longer archive attribute set (some programs don't like
+	it)
+     Fixed a bug that it set badly one flag in large anode tree (it was not
+	destructive)
+1.95 Fixed one buffer leak, that could happen on corrupted filesystem
+     Fixed one bug in allocation in 1.94
+1.96 Added workaround for one bug in OS/2 (HPFS locked up, HPFS386 reported
+	error sometimes when opening directories in PMSHELL)
+     Fixed a possible bitmap race
+     Fixed possible problem on large disks
+     You can now delete open files
+     Fixed a nondestructive race in rename
+1.97 Support for HPFS v3 (on large partitions)
+     Fixed a bug that it didn't allow creation of files > 128M (it should be 2G)
+1.97.1 Changed names of global symbols
+       Fixed a bug when chmoding or chowning root directory
+1.98 Fixed a deadlock when using old_readdir
+     Better directory handling; workaround for "unbalanced tree" bug in OS/2
+1.99 Corrected a possible problem when there's not enough space while deleting
+	file
+     Now it tries to truncate the file if there's not enough space when deleting
+     Removed a lot of redundant code
+2.00 Fixed a bug in rename (it was there since 1.96)
+     Better anti-fragmentation strategy
+2.01 Fixed problem with directory listing over NFS
+     Directory lseek now checks for proper parameters
+     Fixed race-condition in buffer code - it is in all filesystems in Linux;
+        when reading device (cat /dev/hda) while creating files on it, files
+        could be damaged
+2.02 Workaround for bug in breada in Linux. breada could cause accesses beyond
+        end of partition
+2.03 Char, block devices and pipes are correctly created
+     Fixed non-crashing race in unlink (Alexander Viro)
+     Now it works with Japanese version of OS/2
+2.04 Fixed error when ftruncate used to extend file
+2.05 Fixed crash when got mount parameters without =
+     Fixed crash when allocation of anode failed due to full disk
+     Fixed some crashes when block io or inode allocation failed
+2.06 Fixed some crash on corrupted disk structures
+     Better allocation strategy
+     Reschedule points added so that it doesn't lock CPU long time
+     It should work in read-only mode on Warp Server
+2.07 More fixes for Warp Server. Now it really works
+2.08 Creating new files is not so slow on large disks
+     An attempt to sync deleted file does not generate filesystem error
+2.09 Fixed error on extremely fragmented files
+
+
+ vim: set textwidth=80:
diff --git a/Documentation/filesystems/index.rst b/Documentation/filesystems/index.rst
new file mode 100644
index 000000000..46d1b1be3
--- /dev/null
+++ b/Documentation/filesystems/index.rst
@@ -0,0 +1,361 @@
+=====================
+Linux Filesystems API
+=====================
+
+The Linux VFS
+=============
+
+The Filesystem types
+--------------------
+
+.. kernel-doc:: include/linux/fs.h
+   :internal:
+
+The Directory Cache
+-------------------
+
+.. kernel-doc:: fs/dcache.c
+   :export:
+
+.. kernel-doc:: include/linux/dcache.h
+   :internal:
+
+Inode Handling
+--------------
+
+.. kernel-doc:: fs/inode.c
+   :export:
+
+.. kernel-doc:: fs/bad_inode.c
+   :export:
+
+Registration and Superblocks
+----------------------------
+
+.. kernel-doc:: fs/super.c
+   :export:
+
+File Locks
+----------
+
+.. kernel-doc:: fs/locks.c
+   :export:
+
+.. kernel-doc:: fs/locks.c
+   :internal:
+
+Other Functions
+---------------
+
+.. kernel-doc:: fs/mpage.c
+   :export:
+
+.. kernel-doc:: fs/namei.c
+   :export:
+
+.. kernel-doc:: fs/buffer.c
+   :export:
+
+.. kernel-doc:: block/bio.c
+   :export:
+
+.. kernel-doc:: fs/seq_file.c
+   :export:
+
+.. kernel-doc:: fs/filesystems.c
+   :export:
+
+.. kernel-doc:: fs/fs-writeback.c
+   :export:
+
+.. kernel-doc:: fs/block_dev.c
+   :export:
+
+.. kernel-doc:: fs/anon_inodes.c
+   :export:
+
+.. kernel-doc:: fs/attr.c
+   :export:
+
+.. kernel-doc:: fs/d_path.c
+   :export:
+
+.. kernel-doc:: fs/dax.c
+   :export:
+
+.. kernel-doc:: fs/direct-io.c
+   :export:
+
+.. kernel-doc:: fs/file_table.c
+   :export:
+
+.. kernel-doc:: fs/libfs.c
+   :export:
+
+.. kernel-doc:: fs/posix_acl.c
+   :export:
+
+.. kernel-doc:: fs/stat.c
+   :export:
+
+.. kernel-doc:: fs/sync.c
+   :export:
+
+.. kernel-doc:: fs/xattr.c
+   :export:
+
+The proc filesystem
+===================
+
+sysctl interface
+----------------
+
+.. kernel-doc:: kernel/sysctl.c
+   :export:
+
+proc filesystem interface
+-------------------------
+
+.. kernel-doc:: fs/proc/base.c
+   :internal:
+
+Events based on file descriptors
+================================
+
+.. kernel-doc:: fs/eventfd.c
+   :export:
+
+The Filesystem for Exporting Kernel Objects
+===========================================
+
+.. kernel-doc:: fs/sysfs/file.c
+   :export:
+
+.. kernel-doc:: fs/sysfs/symlink.c
+   :export:
+
+The debugfs filesystem
+======================
+
+debugfs interface
+-----------------
+
+.. kernel-doc:: fs/debugfs/inode.c
+   :export:
+
+.. kernel-doc:: fs/debugfs/file.c
+   :export:
+
+The Linux Journalling API
+=========================
+
+Overview
+--------
+
+Details
+~~~~~~~
+
+The journalling layer is easy to use. You need to first of all create a
+journal_t data structure. There are two calls to do this dependent on
+how you decide to allocate the physical media on which the journal
+resides. The :c:func:`jbd2_journal_init_inode` call is for journals stored in
+filesystem inodes, or the :c:func:`jbd2_journal_init_dev` call can be used
+for journal stored on a raw device (in a continuous range of blocks). A
+journal_t is a typedef for a struct pointer, so when you are finally
+finished make sure you call :c:func:`jbd2_journal_destroy` on it to free up
+any used kernel memory.
+
+Once you have got your journal_t object you need to 'mount' or load the
+journal file. The journalling layer expects the space for the journal
+was already allocated and initialized properly by the userspace tools.
+When loading the journal you must call :c:func:`jbd2_journal_load` to process
+journal contents. If the client file system detects the journal contents
+does not need to be processed (or even need not have valid contents), it
+may call :c:func:`jbd2_journal_wipe` to clear the journal contents before
+calling :c:func:`jbd2_journal_load`.
+
+Note that jbd2_journal_wipe(..,0) calls
+:c:func:`jbd2_journal_skip_recovery` for you if it detects any outstanding
+transactions in the journal and similarly :c:func:`jbd2_journal_load` will
+call :c:func:`jbd2_journal_recover` if necessary. I would advise reading
+:c:func:`ext4_load_journal` in fs/ext4/super.c for examples on this stage.
+
+Now you can go ahead and start modifying the underlying filesystem.
+Almost.
+
+You still need to actually journal your filesystem changes, this is done
+by wrapping them into transactions. Additionally you also need to wrap
+the modification of each of the buffers with calls to the journal layer,
+so it knows what the modifications you are actually making are. To do
+this use :c:func:`jbd2_journal_start` which returns a transaction handle.
+
+:c:func:`jbd2_journal_start` and its counterpart :c:func:`jbd2_journal_stop`,
+which indicates the end of a transaction are nestable calls, so you can
+reenter a transaction if necessary, but remember you must call
+:c:func:`jbd2_journal_stop` the same number of times as
+:c:func:`jbd2_journal_start` before the transaction is completed (or more
+accurately leaves the update phase). Ext4/VFS makes use of this feature to
+simplify handling of inode dirtying, quota support, etc.
+
+Inside each transaction you need to wrap the modifications to the
+individual buffers (blocks). Before you start to modify a buffer you
+need to call :c:func:`jbd2_journal_get_create_access()` /
+:c:func:`jbd2_journal_get_write_access()` /
+:c:func:`jbd2_journal_get_undo_access()` as appropriate, this allows the
+journalling layer to copy the unmodified
+data if it needs to. After all the buffer may be part of a previously
+uncommitted transaction. At this point you are at last ready to modify a
+buffer, and once you are have done so you need to call
+:c:func:`jbd2_journal_dirty_metadata`. Or if you've asked for access to a
+buffer you now know is now longer required to be pushed back on the
+device you can call :c:func:`jbd2_journal_forget` in much the same way as you
+might have used :c:func:`bforget` in the past.
+
+A :c:func:`jbd2_journal_flush` may be called at any time to commit and
+checkpoint all your transactions.
+
+Then at umount time , in your :c:func:`put_super` you can then call
+:c:func:`jbd2_journal_destroy` to clean up your in-core journal object.
+
+Unfortunately there a couple of ways the journal layer can cause a
+deadlock. The first thing to note is that each task can only have a
+single outstanding transaction at any one time, remember nothing commits
+until the outermost :c:func:`jbd2_journal_stop`. This means you must complete
+the transaction at the end of each file/inode/address etc. operation you
+perform, so that the journalling system isn't re-entered on another
+journal. Since transactions can't be nested/batched across differing
+journals, and another filesystem other than yours (say ext4) may be
+modified in a later syscall.
+
+The second case to bear in mind is that :c:func:`jbd2_journal_start` can block
+if there isn't enough space in the journal for your transaction (based
+on the passed nblocks param) - when it blocks it merely(!) needs to wait
+for transactions to complete and be committed from other tasks, so
+essentially we are waiting for :c:func:`jbd2_journal_stop`. So to avoid
+deadlocks you must treat :c:func:`jbd2_journal_start` /
+:c:func:`jbd2_journal_stop` as if they were semaphores and include them in
+your semaphore ordering rules to prevent
+deadlocks. Note that :c:func:`jbd2_journal_extend` has similar blocking
+behaviour to :c:func:`jbd2_journal_start` so you can deadlock here just as
+easily as on :c:func:`jbd2_journal_start`.
+
+Try to reserve the right number of blocks the first time. ;-). This will
+be the maximum number of blocks you are going to touch in this
+transaction. I advise having a look at at least ext4_jbd.h to see the
+basis on which ext4 uses to make these decisions.
+
+Another wriggle to watch out for is your on-disk block allocation
+strategy. Why? Because, if you do a delete, you need to ensure you
+haven't reused any of the freed blocks until the transaction freeing
+these blocks commits. If you reused these blocks and crash happens,
+there is no way to restore the contents of the reallocated blocks at the
+end of the last fully committed transaction. One simple way of doing
+this is to mark blocks as free in internal in-memory block allocation
+structures only after the transaction freeing them commits. Ext4 uses
+journal commit callback for this purpose.
+
+With journal commit callbacks you can ask the journalling layer to call
+a callback function when the transaction is finally committed to disk,
+so that you can do some of your own management. You ask the journalling
+layer for calling the callback by simply setting
+``journal->j_commit_callback`` function pointer and that function is
+called after each transaction commit. You can also use
+``transaction->t_private_list`` for attaching entries to a transaction
+that need processing when the transaction commits.
+
+JBD2 also provides a way to block all transaction updates via
+:c:func:`jbd2_journal_lock_updates()` /
+:c:func:`jbd2_journal_unlock_updates()`. Ext4 uses this when it wants a
+window with a clean and stable fs for a moment. E.g.
+
+::
+
+
+        jbd2_journal_lock_updates() //stop new stuff happening..
+        jbd2_journal_flush()        // checkpoint everything.
+        ..do stuff on stable fs
+        jbd2_journal_unlock_updates() // carry on with filesystem use.
+
+The opportunities for abuse and DOS attacks with this should be obvious,
+if you allow unprivileged userspace to trigger codepaths containing
+these calls.
+
+Summary
+~~~~~~~
+
+Using the journal is a matter of wrapping the different context changes,
+being each mount, each modification (transaction) and each changed
+buffer to tell the journalling layer about them.
+
+Data Types
+----------
+
+The journalling layer uses typedefs to 'hide' the concrete definitions
+of the structures used. As a client of the JBD2 layer you can just rely
+on the using the pointer as a magic cookie of some sort. Obviously the
+hiding is not enforced as this is 'C'.
+
+Structures
+~~~~~~~~~~
+
+.. kernel-doc:: include/linux/jbd2.h
+   :internal:
+
+Functions
+---------
+
+The functions here are split into two groups those that affect a journal
+as a whole, and those which are used to manage transactions
+
+Journal Level
+~~~~~~~~~~~~~
+
+.. kernel-doc:: fs/jbd2/journal.c
+   :export:
+
+.. kernel-doc:: fs/jbd2/recovery.c
+   :internal:
+
+Transasction Level
+~~~~~~~~~~~~~~~~~~
+
+.. kernel-doc:: fs/jbd2/transaction.c
+
+See also
+--------
+
+`Journaling the Linux ext2fs Filesystem, LinuxExpo 98, Stephen
+Tweedie <http://kernel.org/pub/linux/kernel/people/sct/ext3/journal-design.ps.gz>`__
+
+`Ext3 Journalling FileSystem, OLS 2000, Dr. Stephen
+Tweedie <http://olstrans.sourceforge.net/release/OLS2000-ext3/OLS2000-ext3.html>`__
+
+splice API
+==========
+
+splice is a method for moving blocks of data around inside the kernel,
+without continually transferring them between the kernel and user space.
+
+.. kernel-doc:: fs/splice.c
+
+pipes API
+=========
+
+Pipe interfaces are all for in-kernel (builtin image) use. They are not
+exported for use by modules.
+
+.. kernel-doc:: include/linux/pipe_fs_i.h
+   :internal:
+
+.. kernel-doc:: fs/pipe.c
+
+Encryption API
+==============
+
+A library which filesystems can hook into to support transparent
+encryption of files and directories.
+
+.. toctree::
+    :maxdepth: 2
+
+    fscrypt
diff --git a/Documentation/filesystems/inotify.txt b/Documentation/filesystems/inotify.txt
new file mode 100644
index 000000000..51f61db78
--- /dev/null
+++ b/Documentation/filesystems/inotify.txt
@@ -0,0 +1,79 @@
+				   inotify
+	    a powerful yet simple file change notification system
+
+
+
+Document started 15 Mar 2005 by Robert Love <rml@novell.com>
+Document updated 4 Jan 2015 by Zhang Zhen <zhenzhang.zhang@huawei.com>
+	--Deleted obsoleted interface, just refer to manpages for user interface.
+
+(i) Rationale
+
+Q: What is the design decision behind not tying the watch to the open fd of
+   the watched object?
+
+A: Watches are associated with an open inotify device, not an open file.
+   This solves the primary problem with dnotify: keeping the file open pins
+   the file and thus, worse, pins the mount.  Dnotify is therefore infeasible
+   for use on a desktop system with removable media as the media cannot be
+   unmounted.  Watching a file should not require that it be open.
+
+Q: What is the design decision behind using an-fd-per-instance as opposed to
+   an fd-per-watch?
+
+A: An fd-per-watch quickly consumes more file descriptors than are allowed,
+   more fd's than are feasible to manage, and more fd's than are optimally
+   select()-able.  Yes, root can bump the per-process fd limit and yes, users
+   can use epoll, but requiring both is a silly and extraneous requirement.
+   A watch consumes less memory than an open file, separating the number
+   spaces is thus sensible.  The current design is what user-space developers
+   want: Users initialize inotify, once, and add n watches, requiring but one
+   fd and no twiddling with fd limits.  Initializing an inotify instance two
+   thousand times is silly.  If we can implement user-space's preferences 
+   cleanly--and we can, the idr layer makes stuff like this trivial--then we 
+   should.
+
+   There are other good arguments.  With a single fd, there is a single
+   item to block on, which is mapped to a single queue of events.  The single
+   fd returns all watch events and also any potential out-of-band data.  If
+   every fd was a separate watch,
+
+   - There would be no way to get event ordering.  Events on file foo and
+     file bar would pop poll() on both fd's, but there would be no way to tell
+     which happened first.  A single queue trivially gives you ordering.  Such
+     ordering is crucial to existing applications such as Beagle.  Imagine
+     "mv a b ; mv b a" events without ordering.
+
+   - We'd have to maintain n fd's and n internal queues with state,
+     versus just one.  It is a lot messier in the kernel.  A single, linear
+     queue is the data structure that makes sense.
+
+   - User-space developers prefer the current API.  The Beagle guys, for
+     example, love it.  Trust me, I asked.  It is not a surprise: Who'd want
+     to manage and block on 1000 fd's via select?
+
+   - No way to get out of band data.
+
+   - 1024 is still too low.  ;-)
+
+   When you talk about designing a file change notification system that
+   scales to 1000s of directories, juggling 1000s of fd's just does not seem
+   the right interface.  It is too heavy.
+
+   Additionally, it _is_ possible to  more than one instance  and
+   juggle more than one queue and thus more than one associated fd.  There
+   need not be a one-fd-per-process mapping; it is one-fd-per-queue and a
+   process can easily want more than one queue.
+
+Q: Why the system call approach?
+
+A: The poor user-space interface is the second biggest problem with dnotify.
+   Signals are a terrible, terrible interface for file notification.  Or for
+   anything, for that matter.  The ideal solution, from all perspectives, is a
+   file descriptor-based one that allows basic file I/O and poll/select.
+   Obtaining the fd and managing the watches could have been done either via a
+   device file or a family of new system calls.  We decided to implement a
+   family of system calls because that is the preferred approach for new kernel
+   interfaces.  The only real difference was whether we wanted to use open(2)
+   and ioctl(2) or a couple of new system calls.  System calls beat ioctls.
+
diff --git a/Documentation/filesystems/isofs.txt b/Documentation/filesystems/isofs.txt
new file mode 100644
index 000000000..ba0a93384
--- /dev/null
+++ b/Documentation/filesystems/isofs.txt
@@ -0,0 +1,48 @@
+Mount options that are the same as for msdos and vfat partitions.
+
+  gid=nnn	All files in the partition will be in group nnn.
+  uid=nnn	All files in the partition will be owned by user id nnn.
+  umask=nnn	The permission mask (see umask(1)) for the partition.
+
+Mount options that are the same as vfat partitions. These are only useful
+when using discs encoded using Microsoft's Joliet extensions.
+  iocharset=name Character set to use for converting from Unicode to
+		ASCII.  Joliet filenames are stored in Unicode format, but
+		Unix for the most part doesn't know how to deal with Unicode.
+		There is also an option of doing UTF-8 translations with the
+		utf8 option.
+  utf8          Encode Unicode names in UTF-8 format. Default is no.
+
+Mount options unique to the isofs filesystem.
+  block=512     Set the block size for the disk to 512 bytes
+  block=1024    Set the block size for the disk to 1024 bytes
+  block=2048    Set the block size for the disk to 2048 bytes
+  check=relaxed Matches filenames with different cases
+  check=strict  Matches only filenames with the exact same case
+  cruft         Try to handle badly formatted CDs.
+  map=off       Do not map non-Rock Ridge filenames to lower case
+  map=normal    Map non-Rock Ridge filenames to lower case
+  map=acorn     As map=normal but also apply Acorn extensions if present
+  mode=xxx      Sets the permissions on files to xxx unless Rock Ridge
+		extensions set the permissions otherwise
+  dmode=xxx     Sets the permissions on directories to xxx unless Rock Ridge
+		extensions set the permissions otherwise
+  overriderockperm Set permissions on files and directories according to
+		'mode' and 'dmode' even though Rock Ridge extensions are
+		present.
+  nojoliet      Ignore Joliet extensions if they are present.
+  norock        Ignore Rock Ridge extensions if they are present.
+  hide		Completely strip hidden files from the file system.
+  showassoc	Show files marked with the 'associated' bit
+  unhide	Deprecated; showing hidden files is now default;
+		If given, it is a synonym for 'showassoc' which will
+		recreate previous unhide behavior
+  session=x     Select number of session on multisession CD
+  sbsector=xxx  Session begins from sector xxx
+
+Recommended documents about ISO 9660 standard are located at:
+http://www.y-adagio.com/
+ftp://ftp.ecma.ch/ecma-st/Ecma-119.pdf
+Quoting from the PDF "This 2nd Edition of Standard ECMA-119 is technically 
+identical with ISO 9660.", so it is a valid and gratis substitute of the
+official ISO specification.
diff --git a/Documentation/filesystems/jfs.txt b/Documentation/filesystems/jfs.txt
new file mode 100644
index 000000000..41fd75799
--- /dev/null
+++ b/Documentation/filesystems/jfs.txt
@@ -0,0 +1,52 @@
+IBM's Journaled File System (JFS) for Linux
+
+JFS Homepage:  http://jfs.sourceforge.net/
+
+The following mount options are supported:
+(*) == default
+
+iocharset=name	Character set to use for converting from Unicode to
+		ASCII.  The default is to do no conversion.  Use
+		iocharset=utf8 for UTF-8 translations.  This requires
+		CONFIG_NLS_UTF8 to be set in the kernel .config file.
+		iocharset=none specifies the default behavior explicitly.
+
+resize=value	Resize the volume to <value> blocks.  JFS only supports
+		growing a volume, not shrinking it.  This option is only
+		valid during a remount, when the volume is mounted
+		read-write.  The resize keyword with no value will grow
+		the volume to the full size of the partition.
+
+nointegrity	Do not write to the journal.  The primary use of this option
+		is to allow for higher performance when restoring a volume
+		from backup media.  The integrity of the volume is not
+		guaranteed if the system abnormally abends.
+
+integrity(*)	Commit metadata changes to the journal.  Use this option to
+		remount a volume where the nointegrity option was
+		previously specified in order to restore normal behavior.
+
+errors=continue		Keep going on a filesystem error.
+errors=remount-ro(*)	Remount the filesystem read-only on an error.
+errors=panic		Panic and halt the machine if an error occurs.
+
+uid=value	Override on-disk uid with specified value
+gid=value	Override on-disk gid with specified value
+umask=value	Override on-disk umask with specified octal value.  For
+		directories, the execute bit will be set if the corresponding
+		read bit is set.
+
+discard=minlen	This enables/disables the use of discard/TRIM commands.
+discard		The discard/TRIM commands are sent to the underlying
+nodiscard(*)	block device when blocks are freed. This is useful for SSD
+		devices and sparse/thinly-provisioned LUNs.  The FITRIM ioctl
+		command is also available together with the nodiscard option.
+		The value of minlen specifies the minimum blockcount, when
+		a TRIM command to the block device is considered useful.
+		When no value is given to the discard option, it defaults to
+		64 blocks, which means 256KiB in JFS.
+		The minlen value of discard overrides the minlen value given
+		on an FITRIM ioctl().
+
+The JFS mailing list can be subscribed to by using the link labeled
+"Mail list Subscribe" at our web page http://jfs.sourceforge.net/
diff --git a/Documentation/filesystems/locks.txt b/Documentation/filesystems/locks.txt
new file mode 100644
index 000000000..5368690f4
--- /dev/null
+++ b/Documentation/filesystems/locks.txt
@@ -0,0 +1,68 @@
+		      File Locking Release Notes
+
+		Andy Walker <andy@lysaker.kvaerner.no>
+
+			    12 May 1997
+
+
+1. What's New?
+--------------
+
+1.1 Broken Flock Emulation
+--------------------------
+
+The old flock(2) emulation in the kernel was swapped for proper BSD
+compatible flock(2) support in the 1.3.x series of kernels. With the
+release of the 2.1.x kernel series, support for the old emulation has
+been totally removed, so that we don't need to carry this baggage
+forever.
+
+This should not cause problems for anybody, since everybody using a
+2.1.x kernel should have updated their C library to a suitable version
+anyway (see the file "Documentation/process/changes.rst".)
+
+1.2 Allow Mixed Locks Again
+---------------------------
+
+1.2.1 Typical Problems - Sendmail
+---------------------------------
+Because sendmail was unable to use the old flock() emulation, many sendmail
+installations use fcntl() instead of flock(). This is true of Slackware 3.0
+for example. This gave rise to some other subtle problems if sendmail was
+configured to rebuild the alias file. Sendmail tried to lock the aliases.dir
+file with fcntl() at the same time as the GDBM routines tried to lock this
+file with flock(). With pre 1.3.96 kernels this could result in deadlocks that,
+over time, or under a very heavy mail load, would eventually cause the kernel
+to lock solid with deadlocked processes.
+
+
+1.2.2 The Solution
+------------------
+The solution I have chosen, after much experimentation and discussion,
+is to make flock() and fcntl() locks oblivious to each other. Both can
+exists, and neither will have any effect on the other.
+
+I wanted the two lock styles to be cooperative, but there were so many
+race and deadlock conditions that the current solution was the only
+practical one. It puts us in the same position as, for example, SunOS
+4.1.x and several other commercial Unices. The only OS's that support
+cooperative flock()/fcntl() are those that emulate flock() using
+fcntl(), with all the problems that implies.
+
+
+1.3 Mandatory Locking As A Mount Option
+---------------------------------------
+
+Mandatory locking, as described in
+'Documentation/filesystems/mandatory-locking.txt' was prior to this release a
+general configuration option that was valid for all mounted filesystems.  This
+had a number of inherent dangers, not the least of which was the ability to
+freeze an NFS server by asking it to read a file for which a mandatory lock
+existed.
+
+From this release of the kernel, mandatory locking can be turned on and off
+on a per-filesystem basis, using the mount options 'mand' and 'nomand'.
+The default is to disallow mandatory locking. The intention is that
+mandatory locking only be enabled on a local filesystem as the specific need
+arises.
+
diff --git a/Documentation/filesystems/mandatory-locking.txt b/Documentation/filesystems/mandatory-locking.txt
new file mode 100644
index 000000000..a251ca331
--- /dev/null
+++ b/Documentation/filesystems/mandatory-locking.txt
@@ -0,0 +1,181 @@
+	Mandatory File Locking For The Linux Operating System
+
+		Andy Walker <andy@lysaker.kvaerner.no>
+
+			   15 April 1996
+		     (Updated September 2007)
+
+0. Why you should avoid mandatory locking
+-----------------------------------------
+
+The Linux implementation is prey to a number of difficult-to-fix race
+conditions which in practice make it not dependable:
+
+	- The write system call checks for a mandatory lock only once
+	  at its start.  It is therefore possible for a lock request to
+	  be granted after this check but before the data is modified.
+	  A process may then see file data change even while a mandatory
+	  lock was held.
+	- Similarly, an exclusive lock may be granted on a file after
+	  the kernel has decided to proceed with a read, but before the
+	  read has actually completed, and the reading process may see
+	  the file data in a state which should not have been visible
+	  to it.
+	- Similar races make the claimed mutual exclusion between lock
+	  and mmap similarly unreliable.
+
+1. What is  mandatory locking?
+------------------------------
+
+Mandatory locking is kernel enforced file locking, as opposed to the more usual
+cooperative file locking used to guarantee sequential access to files among
+processes. File locks are applied using the flock() and fcntl() system calls
+(and the lockf() library routine which is a wrapper around fcntl().) It is
+normally a process' responsibility to check for locks on a file it wishes to
+update, before applying its own lock, updating the file and unlocking it again.
+The most commonly used example of this (and in the case of sendmail, the most
+troublesome) is access to a user's mailbox. The mail user agent and the mail
+transfer agent must guard against updating the mailbox at the same time, and
+prevent reading the mailbox while it is being updated.
+
+In a perfect world all processes would use and honour a cooperative, or
+"advisory" locking scheme. However, the world isn't perfect, and there's
+a lot of poorly written code out there.
+
+In trying to address this problem, the designers of System V UNIX came up
+with a "mandatory" locking scheme, whereby the operating system kernel would
+block attempts by a process to write to a file that another process holds a
+"read" -or- "shared" lock on, and block attempts to both read and write to a 
+file that a process holds a "write " -or- "exclusive" lock on.
+
+The System V mandatory locking scheme was intended to have as little impact as
+possible on existing user code. The scheme is based on marking individual files
+as candidates for mandatory locking, and using the existing fcntl()/lockf()
+interface for applying locks just as if they were normal, advisory locks.
+
+Note 1: In saying "file" in the paragraphs above I am actually not telling
+the whole truth. System V locking is based on fcntl(). The granularity of
+fcntl() is such that it allows the locking of byte ranges in files, in addition
+to entire files, so the mandatory locking rules also have byte level
+granularity.
+
+Note 2: POSIX.1 does not specify any scheme for mandatory locking, despite
+borrowing the fcntl() locking scheme from System V. The mandatory locking
+scheme is defined by the System V Interface Definition (SVID) Version 3.
+
+2. Marking a file for mandatory locking
+---------------------------------------
+
+A file is marked as a candidate for mandatory locking by setting the group-id
+bit in its file mode but removing the group-execute bit. This is an otherwise
+meaningless combination, and was chosen by the System V implementors so as not
+to break existing user programs.
+
+Note that the group-id bit is usually automatically cleared by the kernel when
+a setgid file is written to. This is a security measure. The kernel has been
+modified to recognize the special case of a mandatory lock candidate and to
+refrain from clearing this bit. Similarly the kernel has been modified not
+to run mandatory lock candidates with setgid privileges.
+
+3. Available implementations
+----------------------------
+
+I have considered the implementations of mandatory locking available with
+SunOS 4.1.x, Solaris 2.x and HP-UX 9.x.
+
+Generally I have tried to make the most sense out of the behaviour exhibited
+by these three reference systems. There are many anomalies.
+
+All the reference systems reject all calls to open() for a file on which
+another process has outstanding mandatory locks. This is in direct
+contravention of SVID 3, which states that only calls to open() with the
+O_TRUNC flag set should be rejected. The Linux implementation follows the SVID
+definition, which is the "Right Thing", since only calls with O_TRUNC can
+modify the contents of the file.
+
+HP-UX even disallows open() with O_TRUNC for a file with advisory locks, not
+just mandatory locks. That would appear to contravene POSIX.1.
+
+mmap() is another interesting case. All the operating systems mentioned
+prevent mandatory locks from being applied to an mmap()'ed file, but  HP-UX
+also disallows advisory locks for such a file. SVID actually specifies the
+paranoid HP-UX behaviour.
+
+In my opinion only MAP_SHARED mappings should be immune from locking, and then
+only from mandatory locks - that is what is currently implemented.
+
+SunOS is so hopeless that it doesn't even honour the O_NONBLOCK flag for
+mandatory locks, so reads and writes to locked files always block when they
+should return EAGAIN.
+
+I'm afraid that this is such an esoteric area that the semantics described
+below are just as valid as any others, so long as the main points seem to
+agree. 
+
+4. Semantics
+------------
+
+1. Mandatory locks can only be applied via the fcntl()/lockf() locking
+   interface - in other words the System V/POSIX interface. BSD style
+   locks using flock() never result in a mandatory lock.
+
+2. If a process has locked a region of a file with a mandatory read lock, then
+   other processes are permitted to read from that region. If any of these
+   processes attempts to write to the region it will block until the lock is
+   released, unless the process has opened the file with the O_NONBLOCK
+   flag in which case the system call will return immediately with the error
+   status EAGAIN.
+
+3. If a process has locked a region of a file with a mandatory write lock, all
+   attempts to read or write to that region block until the lock is released,
+   unless a process has opened the file with the O_NONBLOCK flag in which case
+   the system call will return immediately with the error status EAGAIN.
+
+4. Calls to open() with O_TRUNC, or to creat(), on a existing file that has
+   any mandatory locks owned by other processes will be rejected with the
+   error status EAGAIN.
+
+5. Attempts to apply a mandatory lock to a file that is memory mapped and
+   shared (via mmap() with MAP_SHARED) will be rejected with the error status
+   EAGAIN.
+
+6. Attempts to create a shared memory map of a file (via mmap() with MAP_SHARED)
+   that has any mandatory locks in effect will be rejected with the error status
+   EAGAIN.
+
+5. Which system calls are affected?
+-----------------------------------
+
+Those which modify a file's contents, not just the inode. That gives read(),
+write(), readv(), writev(), open(), creat(), mmap(), truncate() and
+ftruncate(). truncate() and ftruncate() are considered to be "write" actions
+for the purposes of mandatory locking.
+
+The affected region is usually defined as stretching from the current position
+for the total number of bytes read or written. For the truncate calls it is
+defined as the bytes of a file removed or added (we must also consider bytes
+added, as a lock can specify just "the whole file", rather than a specific
+range of bytes.)
+
+Note 3: I may have overlooked some system calls that need mandatory lock
+checking in my eagerness to get this code out the door. Please let me know, or
+better still fix the system calls yourself and submit a patch to me or Linus.
+
+6. Warning!
+-----------
+
+Not even root can override a mandatory lock, so runaway processes can wreak
+havoc if they lock crucial files. The way around it is to change the file
+permissions (remove the setgid bit) before trying to read or write to it.
+Of course, that might be a bit tricky if the system is hung :-(
+
+7. The "mand" mount option
+--------------------------
+Mandatory locking is disabled on all filesystems by default, and must be
+administratively enabled by mounting with "-o mand". That mount option
+is only allowed if the mounting task has the CAP_SYS_ADMIN capability.
+
+Since kernel v4.5, it is possible to disable mandatory locking
+altogether by setting CONFIG_MANDATORY_FILE_LOCKING to "n". A kernel
+with this disabled will reject attempts to mount filesystems with the
+"mand" mount option with the error status EPERM.
diff --git a/Documentation/filesystems/nfs/00-INDEX b/Documentation/filesystems/nfs/00-INDEX
new file mode 100644
index 000000000..53f3b596a
--- /dev/null
+++ b/Documentation/filesystems/nfs/00-INDEX
@@ -0,0 +1,26 @@
+00-INDEX
+	- this file (nfs-related documentation).
+Exporting
+	- explanation of how to make filesystems exportable.
+fault_injection.txt
+	- information for using fault injection on the server
+knfsd-stats.txt
+	- statistics which the NFS server makes available to user space.
+nfs.txt
+	- nfs client, and DNS resolution for fs_locations.
+nfs41-server.txt
+	- info on the Linux server implementation of NFSv4 minor version 1.
+nfs-rdma.txt
+	- how to install and setup the Linux NFS/RDMA client and server software
+nfsd-admin-interfaces.txt
+	- Administrative interfaces for nfsd.
+nfsroot.txt
+	- short guide on setting up a diskless box with NFS root filesystem.
+pnfs.txt
+	- short explanation of some of the internals of the pnfs client code
+rpc-cache.txt
+	- introduction to the caching mechanisms in the sunrpc layer.
+idmapper.txt
+	- information for configuring request-keys to be used by idmapper
+rpc-server-gss.txt
+	- Information on GSS authentication support in the NFS Server
diff --git a/Documentation/filesystems/nfs/Exporting b/Documentation/filesystems/nfs/Exporting
new file mode 100644
index 000000000..63889149f
--- /dev/null
+++ b/Documentation/filesystems/nfs/Exporting
@@ -0,0 +1,160 @@
+
+Making Filesystems Exportable
+=============================
+
+Overview
+--------
+
+All filesystem operations require a dentry (or two) as a starting
+point.  Local applications have a reference-counted hold on suitable
+dentries via open file descriptors or cwd/root.  However remote
+applications that access a filesystem via a remote filesystem protocol
+such as NFS may not be able to hold such a reference, and so need a
+different way to refer to a particular dentry.  As the alternative
+form of reference needs to be stable across renames, truncates, and
+server-reboot (among other things, though these tend to be the most
+problematic), there is no simple answer like 'filename'.
+
+The mechanism discussed here allows each filesystem implementation to
+specify how to generate an opaque (outside of the filesystem) byte
+string for any dentry, and how to find an appropriate dentry for any
+given opaque byte string.
+This byte string will be called a "filehandle fragment" as it
+corresponds to part of an NFS filehandle.
+
+A filesystem which supports the mapping between filehandle fragments
+and dentries will be termed "exportable".
+
+
+
+Dcache Issues
+-------------
+
+The dcache normally contains a proper prefix of any given filesystem
+tree.  This means that if any filesystem object is in the dcache, then
+all of the ancestors of that filesystem object are also in the dcache.
+As normal access is by filename this prefix is created naturally and
+maintained easily (by each object maintaining a reference count on
+its parent).
+
+However when objects are included into the dcache by interpreting a
+filehandle fragment, there is no automatic creation of a path prefix
+for the object.  This leads to two related but distinct features of
+the dcache that are not needed for normal filesystem access.
+
+1/ The dcache must sometimes contain objects that are not part of the
+   proper prefix. i.e that are not connected to the root.
+2/ The dcache must be prepared for a newly found (via ->lookup) directory
+   to already have a (non-connected) dentry, and must be able to move
+   that dentry into place (based on the parent and name in the
+   ->lookup).   This is particularly needed for directories as
+   it is a dcache invariant that directories only have one dentry.
+
+To implement these features, the dcache has:
+
+a/ A dentry flag DCACHE_DISCONNECTED which is set on
+   any dentry that might not be part of the proper prefix.
+   This is set when anonymous dentries are created, and cleared when a
+   dentry is noticed to be a child of a dentry which is in the proper
+   prefix.  If the refcount on a dentry with this flag set
+   becomes zero, the dentry is immediately discarded, rather than being
+   kept in the dcache.  If a dentry that is not already in the dcache
+   is repeatedly accessed by filehandle (as NFSD might do), an new dentry
+   will be a allocated for each access, and discarded at the end of
+   the access.
+
+   Note that such a dentry can acquire children, name, ancestors, etc.
+   without losing DCACHE_DISCONNECTED - that flag is only cleared when
+   subtree is successfully reconnected to root.  Until then dentries
+   in such subtree are retained only as long as there are references;
+   refcount reaching zero means immediate eviction, same as for unhashed
+   dentries.  That guarantees that we won't need to hunt them down upon
+   umount.
+
+b/ A primitive for creation of secondary roots - d_obtain_root(inode).
+   Those do _not_ bear DCACHE_DISCONNECTED.  They are placed on the
+   per-superblock list (->s_roots), so they can be located at umount
+   time for eviction purposes.
+
+c/ Helper routines to allocate anonymous dentries, and to help attach
+   loose directory dentries at lookup time. They are:
+    d_obtain_alias(inode) will return a dentry for the given inode.
+      If the inode already has a dentry, one of those is returned.
+      If it doesn't, a new anonymous (IS_ROOT and
+        DCACHE_DISCONNECTED) dentry is allocated and attached.
+      In the case of a directory, care is taken that only one dentry
+      can ever be attached.
+    d_splice_alias(inode, dentry) will introduce a new dentry into the tree;
+      either the passed-in dentry or a preexisting alias for the given inode
+      (such as an anonymous one created by d_obtain_alias), if appropriate.
+      It returns NULL when the passed-in dentry is used, following the calling
+      convention of ->lookup.
+ 
+Filesystem Issues
+-----------------
+
+For a filesystem to be exportable it must:
+ 
+   1/ provide the filehandle fragment routines described below.
+   2/ make sure that d_splice_alias is used rather than d_add
+      when ->lookup finds an inode for a given parent and name.
+
+      If inode is NULL, d_splice_alias(inode, dentry) is equivalent to
+
+		d_add(dentry, inode), NULL
+
+      Similarly, d_splice_alias(ERR_PTR(err), dentry) = ERR_PTR(err)
+
+      Typically the ->lookup routine will simply end with a:
+
+		return d_splice_alias(inode, dentry);
+	}
+
+
+
+  A file system implementation declares that instances of the filesystem
+are exportable by setting the s_export_op field in the struct
+super_block.  This field must point to a "struct export_operations"
+struct which has the following members:
+
+ encode_fh  (optional)
+    Takes a dentry and creates a filehandle fragment which can later be used
+    to find or create a dentry for the same object.  The default
+    implementation creates a filehandle fragment that encodes a 32bit inode
+    and generation number for the inode encoded, and if necessary the
+    same information for the parent.
+
+  fh_to_dentry (mandatory)
+    Given a filehandle fragment, this should find the implied object and
+    create a dentry for it (possibly with d_obtain_alias).
+
+  fh_to_parent (optional but strongly recommended)
+    Given a filehandle fragment, this should find the parent of the
+    implied object and create a dentry for it (possibly with
+    d_obtain_alias).  May fail if the filehandle fragment is too small.
+
+  get_parent (optional but strongly recommended)
+    When given a dentry for a directory, this should return  a dentry for
+    the parent.  Quite possibly the parent dentry will have been allocated
+    by d_alloc_anon.  The default get_parent function just returns an error
+    so any filehandle lookup that requires finding a parent will fail.
+    ->lookup("..") is *not* used as a default as it can leave ".." entries
+    in the dcache which are too messy to work with.
+
+  get_name (optional)
+    When given a parent dentry and a child dentry, this should find a name
+    in the directory identified by the parent dentry, which leads to the
+    object identified by the child dentry.  If no get_name function is
+    supplied, a default implementation is provided which uses vfs_readdir
+    to find potential names, and matches inode numbers to find the correct
+    match.
+
+
+A filehandle fragment consists of an array of 1 or more 4byte words,
+together with a one byte "type".
+The decode_fh routine should not depend on the stated size that is
+passed to it.  This size may be larger than the original filehandle
+generated by encode_fh, in which case it will have been padded with
+nuls.  Rather, the encode_fh routine should choose a "type" which
+indicates the decode_fh how much of the filehandle is valid, and how
+it should be interpreted.
diff --git a/Documentation/filesystems/nfs/fault_injection.txt b/Documentation/filesystems/nfs/fault_injection.txt
new file mode 100644
index 000000000..f3a5b0a8a
--- /dev/null
+++ b/Documentation/filesystems/nfs/fault_injection.txt
@@ -0,0 +1,69 @@
+
+Fault Injection
+===============
+Fault injection is a method for forcing errors that may not normally occur, or
+may be difficult to reproduce.  Forcing these errors in a controlled environment
+can help the developer find and fix bugs before their code is shipped in a
+production system.  Injecting an error on the Linux NFS server will allow us to
+observe how the client reacts and if it manages to recover its state correctly.
+
+NFSD_FAULT_INJECTION must be selected when configuring the kernel to use this
+feature.
+
+
+Using Fault Injection
+=====================
+On the client, mount the fault injection server through NFS v4.0+ and do some
+work over NFS (open files, take locks, ...).
+
+On the server, mount the debugfs filesystem to <debug_dir> and ls
+<debug_dir>/nfsd.  This will show a list of files that will be used for
+injecting faults on the NFS server.  As root, write a number n to the file
+corresponding to the action you want the server to take.  The server will then
+process the first n items it finds.  So if you want to forget 5 locks, echo '5'
+to <debug_dir>/nfsd/forget_locks.  A value of 0 will tell the server to forget
+all corresponding items.  A log message will be created containing the number
+of items forgotten (check dmesg).
+
+Go back to work on the client and check if the client recovered from the error
+correctly.
+
+
+Available Faults
+================
+forget_clients:
+     The NFS server keeps a list of clients that have placed a mount call.  If
+     this list is cleared, the server will have no knowledge of who the client
+     is, forcing the client to reauthenticate with the server.
+
+forget_openowners:
+     The NFS server keeps a list of what files are currently opened and who
+     they were opened by.  Clearing this list will force the client to reopen
+     its files.
+
+forget_locks:
+     The NFS server keeps a list of what files are currently locked in the VFS.
+     Clearing this list will force the client to reclaim its locks (files are
+     unlocked through the VFS as they are cleared from this list).
+
+forget_delegations:
+     A delegation is used to assure the client that a file, or part of a file,
+     has not changed since the delegation was awarded.  Clearing this list will
+     force the client to reacquire its delegation before accessing the file
+     again.
+
+recall_delegations:
+     Delegations can be recalled by the server when another client attempts to
+     access a file.  This test will notify the client that its delegation has
+     been revoked, forcing the client to reacquire the delegation before using
+     the file again.
+
+
+tools/nfs/inject_faults.sh script
+=================================
+This script has been created to ease the fault injection process.  This script
+will detect the mounted debugfs directory and write to the files located there
+based on the arguments passed by the user.  For example, running
+`inject_faults.sh forget_locks 1` as root will instruct the server to forget
+one lock.  Running `inject_faults forget_locks` will instruct the server to
+forgetall locks.
diff --git a/Documentation/filesystems/nfs/idmapper.txt b/Documentation/filesystems/nfs/idmapper.txt
new file mode 100644
index 000000000..b86831acd
--- /dev/null
+++ b/Documentation/filesystems/nfs/idmapper.txt
@@ -0,0 +1,75 @@
+
+=========
+ID Mapper
+=========
+Id mapper is used by NFS to translate user and group ids into names, and to
+translate user and group names into ids.  Part of this translation involves
+performing an upcall to userspace to request the information.  There are two
+ways NFS could obtain this information: placing a call to /sbin/request-key
+or by placing a call to the rpc.idmap daemon.
+
+NFS will attempt to call /sbin/request-key first.  If this succeeds, the
+result will be cached using the generic request-key cache.  This call should
+only fail if /etc/request-key.conf is not configured for the id_resolver key
+type, see the "Configuring" section below if you wish to use the request-key
+method.
+
+If the call to /sbin/request-key fails (if /etc/request-key.conf is not
+configured with the id_resolver key type), then the idmapper will ask the
+legacy rpc.idmap daemon for the id mapping.  This result will be stored
+in a custom NFS idmap cache.
+
+
+===========
+Configuring
+===========
+The file /etc/request-key.conf will need to be modified so /sbin/request-key can
+direct the upcall.  The following line should be added:
+
+#OP	TYPE	DESCRIPTION	CALLOUT INFO	PROGRAM ARG1 ARG2 ARG3 ...
+#======	=======	===============	===============	===============================
+create	id_resolver	*	*		/usr/sbin/nfs.idmap %k %d 600
+
+This will direct all id_resolver requests to the program /usr/sbin/nfs.idmap.
+The last parameter, 600, defines how many seconds into the future the key will
+expire.  This parameter is optional for /usr/sbin/nfs.idmap.  When the timeout
+is not specified, nfs.idmap will default to 600 seconds.
+
+id mapper uses for key descriptions:
+	  uid:  Find the UID for the given user
+	  gid:  Find the GID for the given group
+	 user:  Find the user  name for the given UID
+	group:  Find the group name for the given GID
+
+You can handle any of these individually, rather than using the generic upcall
+program.  If you would like to use your own program for a uid lookup then you
+would edit your request-key.conf so it look similar to this:
+
+#OP	TYPE	DESCRIPTION	CALLOUT INFO	PROGRAM ARG1 ARG2 ARG3 ...
+#======	=======	===============	===============	===============================
+create	id_resolver	uid:*	*		/some/other/program %k %d 600
+create	id_resolver	*	*		/usr/sbin/nfs.idmap %k %d 600
+
+Notice that the new line was added above the line for the generic program.
+request-key will find the first matching line and corresponding program.  In
+this case, /some/other/program will handle all uid lookups and
+/usr/sbin/nfs.idmap will handle gid, user, and group lookups.
+
+See <file:Documentation/security/keys/request-key.rst> for more information
+about the request-key function.
+
+
+=========
+nfs.idmap
+=========
+nfs.idmap is designed to be called by request-key, and should not be run "by
+hand".  This program takes two arguments, a serialized key and a key
+description.  The serialized key is first converted into a key_serial_t, and
+then passed as an argument to keyctl_instantiate (both are part of keyutils.h).
+
+The actual lookups are performed by functions found in nfsidmap.h.  nfs.idmap
+determines the correct function to call by looking at the first part of the
+description string.  For example, a uid lookup description will appear as
+"uid:user@domain".
+
+nfs.idmap will return 0 if the key was instantiated, and non-zero otherwise.
diff --git a/Documentation/filesystems/nfs/knfsd-stats.txt b/Documentation/filesystems/nfs/knfsd-stats.txt
new file mode 100644
index 000000000..1a5d82180
--- /dev/null
+++ b/Documentation/filesystems/nfs/knfsd-stats.txt
@@ -0,0 +1,123 @@
+
+Kernel NFS Server Statistics
+============================
+
+This document describes the format and semantics of the statistics
+which the kernel NFS server makes available to userspace.  These
+statistics are available in several text form pseudo files, each of
+which is described separately below.
+
+In most cases you don't need to know these formats, as the nfsstat(8)
+program from the nfs-utils distribution provides a helpful command-line
+interface for extracting and printing them.
+
+All the files described here are formatted as a sequence of text lines,
+separated by newline '\n' characters.  Lines beginning with a hash
+'#' character are comments intended for humans and should be ignored
+by parsing routines.  All other lines contain a sequence of fields
+separated by whitespace.
+
+/proc/fs/nfsd/pool_stats
+------------------------
+
+This file is available in kernels from 2.6.30 onwards, if the
+/proc/fs/nfsd filesystem is mounted (it almost always should be).
+
+The first line is a comment which describes the fields present in
+all the other lines.  The other lines present the following data as
+a sequence of unsigned decimal numeric fields.  One line is shown
+for each NFS thread pool.
+
+All counters are 64 bits wide and wrap naturally.  There is no way
+to zero these counters, instead applications should do their own
+rate conversion.
+
+pool
+	The id number of the NFS thread pool to which this line applies.
+	This number does not change.
+
+	Thread pool ids are a contiguous set of small integers starting
+	at zero.  The maximum value depends on the thread pool mode, but
+	currently cannot be larger than the number of CPUs in the system.
+	Note that in the default case there will be a single thread pool
+	which contains all the nfsd threads and all the CPUs in the system,
+	and thus this file will have a single line with a pool id of "0".
+
+packets-arrived
+	Counts how many NFS packets have arrived.  More precisely, this
+	is the number of times that the network stack has notified the
+	sunrpc server layer that new data may be available on a transport
+	(e.g. an NFS or UDP socket or an NFS/RDMA endpoint).
+
+	Depending on the NFS workload patterns and various network stack
+	effects (such as Large Receive Offload) which can combine packets
+	on the wire, this may be either more or less than the number
+	of NFS calls received (which statistic is available elsewhere).
+	However this is a more accurate and less workload-dependent measure
+	of how much CPU load is being placed on the sunrpc server layer
+	due to NFS network traffic.
+
+sockets-enqueued
+	Counts how many times an NFS transport is enqueued to wait for
+	an nfsd thread to service it, i.e. no nfsd thread was considered
+	available.
+
+	The circumstance this statistic tracks indicates that there was NFS
+	network-facing work to be done but it couldn't be done immediately,
+	thus introducing a small delay in servicing NFS calls.  The ideal
+	rate of change for this counter is zero; significantly non-zero
+	values may indicate a performance limitation.
+
+	This can happen because there are too few nfsd threads in the thread
+	pool for the NFS workload (the workload is thread-limited), in which
+	case configuring more nfsd threads will probably improve the
+	performance of the NFS workload.
+
+threads-woken
+	Counts how many times an idle nfsd thread is woken to try to
+	receive some data from an NFS transport.
+
+	This statistic tracks the circumstance where incoming
+	network-facing NFS work is being handled quickly, which is a good
+	thing.  The ideal rate of change for this counter will be close
+	to but less than the rate of change of the packets-arrived counter.
+
+threads-timedout
+	Counts how many times an nfsd thread triggered an idle timeout,
+	i.e. was not woken to handle any incoming network packets for
+	some time.
+
+	This statistic counts a circumstance where there are more nfsd
+	threads configured than can be used by the NFS workload.  This is
+	a clue that the number of nfsd threads can be reduced without
+	affecting performance.  Unfortunately, it's only a clue and not
+	a strong indication, for a couple of reasons:
+
+	 - Currently the rate at which the counter is incremented is quite
+	   slow; the idle timeout is 60 minutes.  Unless the NFS workload
+	   remains constant for hours at a time, this counter is unlikely
+	   to be providing information that is still useful.
+
+	 - It is usually a wise policy to provide some slack,
+	   i.e. configure a few more nfsds than are currently needed,
+	   to allow for future spikes in load.
+
+
+Note that incoming packets on NFS transports will be dealt with in
+one of three ways.  An nfsd thread can be woken (threads-woken counts
+this case), or the transport can be enqueued for later attention
+(sockets-enqueued counts this case), or the packet can be temporarily
+deferred because the transport is currently being used by an nfsd
+thread.  This last case is not very interesting and is not explicitly
+counted, but can be inferred from the other counters thus:
+
+packets-deferred = packets-arrived - ( sockets-enqueued + threads-woken )
+
+
+More
+----
+Descriptions of the other statistics file should go here.
+
+
+Greg Banks <gnb@sgi.com>
+26 Mar 2009
diff --git a/Documentation/filesystems/nfs/nfs-rdma.txt b/Documentation/filesystems/nfs/nfs-rdma.txt
new file mode 100644
index 000000000..22dc0dd68
--- /dev/null
+++ b/Documentation/filesystems/nfs/nfs-rdma.txt
@@ -0,0 +1,274 @@
+################################################################################
+#									       #
+#				NFS/RDMA README				       #
+#									       #
+################################################################################
+
+ Author: NetApp and Open Grid Computing
+ Date: May 29, 2008
+
+Table of Contents
+~~~~~~~~~~~~~~~~~
+ - Overview
+ - Getting Help
+ - Installation
+ - Check RDMA and NFS Setup
+ - NFS/RDMA Setup
+
+Overview
+~~~~~~~~
+
+  This document describes how to install and setup the Linux NFS/RDMA client
+  and server software.
+
+  The NFS/RDMA client was first included in Linux 2.6.24. The NFS/RDMA server
+  was first included in the following release, Linux 2.6.25.
+
+  In our testing, we have obtained excellent performance results (full 10Gbit
+  wire bandwidth at minimal client CPU) under many workloads. The code passes
+  the full Connectathon test suite and operates over both Infiniband and iWARP
+  RDMA adapters.
+
+Getting Help
+~~~~~~~~~~~~
+
+  If you get stuck, you can ask questions on the
+
+                nfs-rdma-devel@lists.sourceforge.net
+
+  mailing list.
+
+Installation
+~~~~~~~~~~~~
+
+  These instructions are a step by step guide to building a machine for
+  use with NFS/RDMA.
+
+  - Install an RDMA device
+
+    Any device supported by the drivers in drivers/infiniband/hw is acceptable.
+
+    Testing has been performed using several Mellanox-based IB cards, the
+    Ammasso AMS1100 iWARP adapter, and the Chelsio cxgb3 iWARP adapter.
+
+  - Install a Linux distribution and tools
+
+    The first kernel release to contain both the NFS/RDMA client and server was
+    Linux 2.6.25  Therefore, a distribution compatible with this and subsequent
+    Linux kernel release should be installed.
+
+    The procedures described in this document have been tested with
+    distributions from Red Hat's Fedora Project (http://fedora.redhat.com/).
+
+  - Install nfs-utils-1.1.2 or greater on the client
+
+    An NFS/RDMA mount point can be obtained by using the mount.nfs command in
+    nfs-utils-1.1.2 or greater (nfs-utils-1.1.1 was the first nfs-utils
+    version with support for NFS/RDMA mounts, but for various reasons we
+    recommend using nfs-utils-1.1.2 or greater). To see which version of
+    mount.nfs you are using, type:
+
+    $ /sbin/mount.nfs -V
+
+    If the version is less than 1.1.2 or the command does not exist,
+    you should install the latest version of nfs-utils.
+
+    Download the latest package from:
+
+    http://www.kernel.org/pub/linux/utils/nfs
+
+    Uncompress the package and follow the installation instructions.
+
+    If you will not need the idmapper and gssd executables (you do not need
+    these to create an NFS/RDMA enabled mount command), the installation
+    process can be simplified by disabling these features when running
+    configure:
+
+    $ ./configure --disable-gss --disable-nfsv4
+
+    To build nfs-utils you will need the tcp_wrappers package installed. For
+    more information on this see the package's README and INSTALL files.
+
+    After building the nfs-utils package, there will be a mount.nfs binary in
+    the utils/mount directory. This binary can be used to initiate NFS v2, v3,
+    or v4 mounts. To initiate a v4 mount, the binary must be called
+    mount.nfs4.  The standard technique is to create a symlink called
+    mount.nfs4 to mount.nfs.
+
+    This mount.nfs binary should be installed at /sbin/mount.nfs as follows:
+
+    $ sudo cp utils/mount/mount.nfs /sbin/mount.nfs
+
+    In this location, mount.nfs will be invoked automatically for NFS mounts
+    by the system mount command.
+
+    NOTE: mount.nfs and therefore nfs-utils-1.1.2 or greater is only needed
+    on the NFS client machine. You do not need this specific version of
+    nfs-utils on the server. Furthermore, only the mount.nfs command from
+    nfs-utils-1.1.2 is needed on the client.
+
+  - Install a Linux kernel with NFS/RDMA
+
+    The NFS/RDMA client and server are both included in the mainline Linux
+    kernel version 2.6.25 and later. This and other versions of the Linux
+    kernel can be found at:
+
+    https://www.kernel.org/pub/linux/kernel/
+
+    Download the sources and place them in an appropriate location.
+
+  - Configure the RDMA stack
+
+    Make sure your kernel configuration has RDMA support enabled. Under
+    Device Drivers -> InfiniBand support, update the kernel configuration
+    to enable InfiniBand support [NOTE: the option name is misleading. Enabling
+    InfiniBand support is required for all RDMA devices (IB, iWARP, etc.)].
+
+    Enable the appropriate IB HCA support (mlx4, mthca, ehca, ipath, etc.) or
+    iWARP adapter support (amso, cxgb3, etc.).
+
+    If you are using InfiniBand, be sure to enable IP-over-InfiniBand support.
+
+  - Configure the NFS client and server
+
+    Your kernel configuration must also have NFS file system support and/or
+    NFS server support enabled. These and other NFS related configuration
+    options can be found under File Systems -> Network File Systems.
+
+  - Build, install, reboot
+
+    The NFS/RDMA code will be enabled automatically if NFS and RDMA
+    are turned on. The NFS/RDMA client and server are configured via the hidden
+    SUNRPC_XPRT_RDMA config option that depends on SUNRPC and INFINIBAND. The
+    value of SUNRPC_XPRT_RDMA will be:
+
+     - N if either SUNRPC or INFINIBAND are N, in this case the NFS/RDMA client
+       and server will not be built
+     - M if both SUNRPC and INFINIBAND are on (M or Y) and at least one is M,
+       in this case the NFS/RDMA client and server will be built as modules
+     - Y if both SUNRPC and INFINIBAND are Y, in this case the NFS/RDMA client
+       and server will be built into the kernel
+
+    Therefore, if you have followed the steps above and turned no NFS and RDMA,
+    the NFS/RDMA client and server will be built.
+
+    Build a new kernel, install it, boot it.
+
+Check RDMA and NFS Setup
+~~~~~~~~~~~~~~~~~~~~~~~~
+
+    Before configuring the NFS/RDMA software, it is a good idea to test
+    your new kernel to ensure that the kernel is working correctly.
+    In particular, it is a good idea to verify that the RDMA stack
+    is functioning as expected and standard NFS over TCP/IP and/or UDP/IP
+    is working properly.
+
+  - Check RDMA Setup
+
+    If you built the RDMA components as modules, load them at
+    this time. For example, if you are using a Mellanox Tavor/Sinai/Arbel
+    card:
+
+    $ modprobe ib_mthca
+    $ modprobe ib_ipoib
+
+    If you are using InfiniBand, make sure there is a Subnet Manager (SM)
+    running on the network. If your IB switch has an embedded SM, you can
+    use it. Otherwise, you will need to run an SM, such as OpenSM, on one
+    of your end nodes.
+
+    If an SM is running on your network, you should see the following:
+
+    $ cat /sys/class/infiniband/driverX/ports/1/state
+    4: ACTIVE
+
+    where driverX is mthca0, ipath5, ehca3, etc.
+
+    To further test the InfiniBand software stack, use IPoIB (this
+    assumes you have two IB hosts named host1 and host2):
+
+    host1$ ip link set dev ib0 up
+    host1$ ip address add dev ib0 a.b.c.x
+    host2$ ip link set dev ib0 up
+    host2$ ip address add dev ib0 a.b.c.y
+    host1$ ping a.b.c.y
+    host2$ ping a.b.c.x
+
+    For other device types, follow the appropriate procedures.
+
+  - Check NFS Setup
+
+    For the NFS components enabled above (client and/or server),
+    test their functionality over standard Ethernet using TCP/IP or UDP/IP.
+
+NFS/RDMA Setup
+~~~~~~~~~~~~~~
+
+  We recommend that you use two machines, one to act as the client and
+  one to act as the server.
+
+  One time configuration:
+
+  - On the server system, configure the /etc/exports file and
+    start the NFS/RDMA server.
+
+    Exports entries with the following formats have been tested:
+
+    /vol0   192.168.0.47(fsid=0,rw,async,insecure,no_root_squash)
+    /vol0   192.168.0.0/255.255.255.0(fsid=0,rw,async,insecure,no_root_squash)
+
+    The IP address(es) is(are) the client's IPoIB address for an InfiniBand
+    HCA or the client's iWARP address(es) for an RNIC.
+
+    NOTE: The "insecure" option must be used because the NFS/RDMA client does
+    not use a reserved port.
+
+ Each time a machine boots:
+
+  - Load and configure the RDMA drivers
+
+    For InfiniBand using a Mellanox adapter:
+
+    $ modprobe ib_mthca
+    $ modprobe ib_ipoib
+    $ ip li set dev ib0 up
+    $ ip addr add dev ib0 a.b.c.d
+
+    NOTE: use unique addresses for the client and server
+
+  - Start the NFS server
+
+    If the NFS/RDMA server was built as a module (CONFIG_SUNRPC_XPRT_RDMA=m in
+    kernel config), load the RDMA transport module:
+
+    $ modprobe svcrdma
+
+    Regardless of how the server was built (module or built-in), start the
+    server:
+
+    $ /etc/init.d/nfs start
+
+    or
+
+    $ service nfs start
+
+    Instruct the server to listen on the RDMA transport:
+
+    $ echo rdma 20049 > /proc/fs/nfsd/portlist
+
+  - On the client system
+
+    If the NFS/RDMA client was built as a module (CONFIG_SUNRPC_XPRT_RDMA=m in
+    kernel config), load the RDMA client module:
+
+    $ modprobe xprtrdma.ko
+
+    Regardless of how the client was built (module or built-in), use this
+    command to mount the NFS/RDMA server:
+
+    $ mount -o rdma,port=20049 <IPoIB-server-name-or-address>:/<export> /mnt
+
+    To verify that the mount is using RDMA, run "cat /proc/mounts" and check
+    the "proto" field for the given mount.
+
+  Congratulations! You're using NFS/RDMA!
diff --git a/Documentation/filesystems/nfs/nfs.txt b/Documentation/filesystems/nfs/nfs.txt
new file mode 100644
index 000000000..f2571c8be
--- /dev/null
+++ b/Documentation/filesystems/nfs/nfs.txt
@@ -0,0 +1,136 @@
+
+The NFS client
+==============
+
+The NFS version 2 protocol was first documented in RFC1094 (March 1989).
+Since then two more major releases of NFS have been published, with NFSv3
+being documented in RFC1813 (June 1995), and NFSv4 in RFC3530 (April
+2003).
+
+The Linux NFS client currently supports all the above published versions,
+and work is in progress on adding support for minor version 1 of the NFSv4
+protocol.
+
+The purpose of this document is to provide information on some of the
+special features of the NFS client that can be configured by system
+administrators.
+
+
+The nfs4_unique_id parameter
+============================
+
+NFSv4 requires clients to identify themselves to servers with a unique
+string.  File open and lock state shared between one client and one server
+is associated with this identity.  To support robust NFSv4 state recovery
+and transparent state migration, this identity string must not change
+across client reboots.
+
+Without any other intervention, the Linux client uses a string that contains
+the local system's node name.  System administrators, however, often do not
+take care to ensure that node names are fully qualified and do not change
+over the lifetime of a client system.  Node names can have other
+administrative requirements that require particular behavior that does not
+work well as part of an nfs_client_id4 string.
+
+The nfs.nfs4_unique_id boot parameter specifies a unique string that can be
+used instead of a system's node name when an NFS client identifies itself to
+a server.  Thus, if the system's node name is not unique, or it changes, its
+nfs.nfs4_unique_id stays the same, preventing collision with other clients
+or loss of state during NFS reboot recovery or transparent state migration.
+
+The nfs.nfs4_unique_id string is typically a UUID, though it can contain
+anything that is believed to be unique across all NFS clients.  An
+nfs4_unique_id string should be chosen when a client system is installed,
+just as a system's root file system gets a fresh UUID in its label at
+install time.
+
+The string should remain fixed for the lifetime of the client.  It can be
+changed safely if care is taken that the client shuts down cleanly and all
+outstanding NFSv4 state has expired, to prevent loss of NFSv4 state.
+
+This string can be stored in an NFS client's grub.conf, or it can be provided
+via a net boot facility such as PXE.  It may also be specified as an nfs.ko
+module parameter.  Specifying a uniquifier string is not support for NFS
+clients running in containers.
+
+
+The DNS resolver
+================
+
+NFSv4 allows for one server to refer the NFS client to data that has been
+migrated onto another server by means of the special "fs_locations"
+attribute. See
+	http://tools.ietf.org/html/rfc3530#section-6
+and
+	http://tools.ietf.org/html/draft-ietf-nfsv4-referrals-00
+
+The fs_locations information can take the form of either an ip address and
+a path, or a DNS hostname and a path. The latter requires the NFS client to
+do a DNS lookup in order to mount the new volume, and hence the need for an
+upcall to allow userland to provide this service.
+
+Assuming that the user has the 'rpc_pipefs' filesystem mounted in the usual
+/var/lib/nfs/rpc_pipefs, the upcall consists of the following steps:
+
+   (1) The process checks the dns_resolve cache to see if it contains a
+       valid entry. If so, it returns that entry and exits.
+
+   (2) If no valid entry exists, the helper script '/sbin/nfs_cache_getent'
+       (may be changed using the 'nfs.cache_getent' kernel boot parameter)
+       is run, with two arguments:
+		- the cache name, "dns_resolve"
+		- the hostname to resolve
+
+   (3) After looking up the corresponding ip address, the helper script
+       writes the result into the rpc_pipefs pseudo-file
+       '/var/lib/nfs/rpc_pipefs/cache/dns_resolve/channel'
+       in the following (text) format:
+
+		"<ip address> <hostname> <ttl>\n"
+
+       Where <ip address> is in the usual IPv4 (123.456.78.90) or IPv6
+       (ffee:ddcc:bbaa:9988:7766:5544:3322:1100, ffee::1100, ...) format.
+       <hostname> is identical to the second argument of the helper
+       script, and <ttl> is the 'time to live' of this cache entry (in
+       units of seconds).
+
+       Note: If <ip address> is invalid, say the string "0", then a negative
+       entry is created, which will cause the kernel to treat the hostname
+       as having no valid DNS translation.
+
+
+
+
+A basic sample /sbin/nfs_cache_getent
+=====================================
+
+#!/bin/bash
+#
+ttl=600
+#
+cut=/usr/bin/cut
+getent=/usr/bin/getent
+rpc_pipefs=/var/lib/nfs/rpc_pipefs
+#
+die()
+{
+	echo "Usage: $0 cache_name entry_name"
+	exit 1
+}
+
+[ $# -lt 2 ] && die
+cachename="$1"
+cache_path=${rpc_pipefs}/cache/${cachename}/channel
+
+case "${cachename}" in
+	dns_resolve)
+		name="$2"
+		result="$(${getent} hosts ${name} | ${cut} -f1 -d\ )"
+		[ -z "${result}" ] && result="0"
+		;;
+	*)
+		die
+		;;
+esac
+echo "${result} ${name} ${ttl}" >${cache_path}
+
diff --git a/Documentation/filesystems/nfs/nfs41-server.txt b/Documentation/filesystems/nfs/nfs41-server.txt
new file mode 100644
index 000000000..682a59fab
--- /dev/null
+++ b/Documentation/filesystems/nfs/nfs41-server.txt
@@ -0,0 +1,173 @@
+NFSv4.1 Server Implementation
+
+Server support for minorversion 1 can be controlled using the
+/proc/fs/nfsd/versions control file.  The string output returned
+by reading this file will contain either "+4.1" or "-4.1"
+correspondingly.
+
+Currently, server support for minorversion 1 is enabled by default.
+It can be disabled at run time by writing the string "-4.1" to
+the /proc/fs/nfsd/versions control file.  Note that to write this
+control file, the nfsd service must be taken down.  You can use rpc.nfsd
+for this; see rpc.nfsd(8).
+
+(Warning: older servers will interpret "+4.1" and "-4.1" as "+4" and
+"-4", respectively.  Therefore, code meant to work on both new and old
+kernels must turn 4.1 on or off *before* turning support for version 4
+on or off; rpc.nfsd does this correctly.)
+
+The NFSv4 minorversion 1 (NFSv4.1) implementation in nfsd is based
+on RFC 5661.
+
+From the many new features in NFSv4.1 the current implementation
+focuses on the mandatory-to-implement NFSv4.1 Sessions, providing
+"exactly once" semantics and better control and throttling of the
+resources allocated for each client.
+
+The table below, taken from the NFSv4.1 document, lists
+the operations that are mandatory to implement (REQ), optional
+(OPT), and NFSv4.0 operations that are required not to implement (MNI)
+in minor version 1.  The first column indicates the operations that
+are not supported yet by the linux server implementation.
+
+The OPTIONAL features identified and their abbreviations are as follows:
+	pNFS	Parallel NFS
+	FDELG	File Delegations
+	DDELG	Directory Delegations
+
+The following abbreviations indicate the linux server implementation status.
+	I	Implemented NFSv4.1 operations.
+	NS	Not Supported.
+	NS*	Unimplemented optional feature.
+
+Operations
+
+   +----------------------+------------+--------------+----------------+
+   | Operation            | REQ, REC,  | Feature      | Definition     |
+   |                      | OPT, or    | (REQ, REC,   |                |
+   |                      | MNI        | or OPT)      |                |
+   +----------------------+------------+--------------+----------------+
+   | ACCESS               | REQ        |              | Section 18.1   |
+I  | BACKCHANNEL_CTL      | REQ        |              | Section 18.33  |
+I  | BIND_CONN_TO_SESSION | REQ        |              | Section 18.34  |
+   | CLOSE                | REQ        |              | Section 18.2   |
+   | COMMIT               | REQ        |              | Section 18.3   |
+   | CREATE               | REQ        |              | Section 18.4   |
+I  | CREATE_SESSION       | REQ        |              | Section 18.36  |
+NS*| DELEGPURGE           | OPT        | FDELG (REQ)  | Section 18.5   |
+   | DELEGRETURN          | OPT        | FDELG,       | Section 18.6   |
+   |                      |            | DDELG, pNFS  |                |
+   |                      |            | (REQ)        |                |
+I  | DESTROY_CLIENTID     | REQ        |              | Section 18.50  |
+I  | DESTROY_SESSION      | REQ        |              | Section 18.37  |
+I  | EXCHANGE_ID          | REQ        |              | Section 18.35  |
+I  | FREE_STATEID         | REQ        |              | Section 18.38  |
+   | GETATTR              | REQ        |              | Section 18.7   |
+I  | GETDEVICEINFO        | OPT        | pNFS (REQ)   | Section 18.40  |
+NS*| GETDEVICELIST        | OPT        | pNFS (OPT)   | Section 18.41  |
+   | GETFH                | REQ        |              | Section 18.8   |
+NS*| GET_DIR_DELEGATION   | OPT        | DDELG (REQ)  | Section 18.39  |
+I  | LAYOUTCOMMIT         | OPT        | pNFS (REQ)   | Section 18.42  |
+I  | LAYOUTGET            | OPT        | pNFS (REQ)   | Section 18.43  |
+I  | LAYOUTRETURN         | OPT        | pNFS (REQ)   | Section 18.44  |
+   | LINK                 | OPT        |              | Section 18.9   |
+   | LOCK                 | REQ        |              | Section 18.10  |
+   | LOCKT                | REQ        |              | Section 18.11  |
+   | LOCKU                | REQ        |              | Section 18.12  |
+   | LOOKUP               | REQ        |              | Section 18.13  |
+   | LOOKUPP              | REQ        |              | Section 18.14  |
+   | NVERIFY              | REQ        |              | Section 18.15  |
+   | OPEN                 | REQ        |              | Section 18.16  |
+NS*| OPENATTR             | OPT        |              | Section 18.17  |
+   | OPEN_CONFIRM         | MNI        |              | N/A            |
+   | OPEN_DOWNGRADE       | REQ        |              | Section 18.18  |
+   | PUTFH                | REQ        |              | Section 18.19  |
+   | PUTPUBFH             | REQ        |              | Section 18.20  |
+   | PUTROOTFH            | REQ        |              | Section 18.21  |
+   | READ                 | REQ        |              | Section 18.22  |
+   | READDIR              | REQ        |              | Section 18.23  |
+   | READLINK             | OPT        |              | Section 18.24  |
+   | RECLAIM_COMPLETE     | REQ        |              | Section 18.51  |
+   | RELEASE_LOCKOWNER    | MNI        |              | N/A            |
+   | REMOVE               | REQ        |              | Section 18.25  |
+   | RENAME               | REQ        |              | Section 18.26  |
+   | RENEW                | MNI        |              | N/A            |
+   | RESTOREFH            | REQ        |              | Section 18.27  |
+   | SAVEFH               | REQ        |              | Section 18.28  |
+   | SECINFO              | REQ        |              | Section 18.29  |
+I  | SECINFO_NO_NAME      | REC        | pNFS files   | Section 18.45, |
+   |                      |            | layout (REQ) | Section 13.12  |
+I  | SEQUENCE             | REQ        |              | Section 18.46  |
+   | SETATTR              | REQ        |              | Section 18.30  |
+   | SETCLIENTID          | MNI        |              | N/A            |
+   | SETCLIENTID_CONFIRM  | MNI        |              | N/A            |
+NS | SET_SSV              | REQ        |              | Section 18.47  |
+I  | TEST_STATEID         | REQ        |              | Section 18.48  |
+   | VERIFY               | REQ        |              | Section 18.31  |
+NS*| WANT_DELEGATION      | OPT        | FDELG (OPT)  | Section 18.49  |
+   | WRITE                | REQ        |              | Section 18.32  |
+
+Callback Operations
+
+   +-------------------------+-----------+-------------+---------------+
+   | Operation               | REQ, REC, | Feature     | Definition    |
+   |                         | OPT, or   | (REQ, REC,  |               |
+   |                         | MNI       | or OPT)     |               |
+   +-------------------------+-----------+-------------+---------------+
+   | CB_GETATTR              | OPT       | FDELG (REQ) | Section 20.1  |
+I  | CB_LAYOUTRECALL         | OPT       | pNFS (REQ)  | Section 20.3  |
+NS*| CB_NOTIFY               | OPT       | DDELG (REQ) | Section 20.4  |
+NS*| CB_NOTIFY_DEVICEID      | OPT       | pNFS (OPT)  | Section 20.12 |
+NS*| CB_NOTIFY_LOCK          | OPT       |             | Section 20.11 |
+NS*| CB_PUSH_DELEG           | OPT       | FDELG (OPT) | Section 20.5  |
+   | CB_RECALL               | OPT       | FDELG,      | Section 20.2  |
+   |                         |           | DDELG, pNFS |               |
+   |                         |           | (REQ)       |               |
+NS*| CB_RECALL_ANY           | OPT       | FDELG,      | Section 20.6  |
+   |                         |           | DDELG, pNFS |               |
+   |                         |           | (REQ)       |               |
+NS | CB_RECALL_SLOT          | REQ       |             | Section 20.8  |
+NS*| CB_RECALLABLE_OBJ_AVAIL | OPT       | DDELG, pNFS | Section 20.7  |
+   |                         |           | (REQ)       |               |
+I  | CB_SEQUENCE             | OPT       | FDELG,      | Section 20.9  |
+   |                         |           | DDELG, pNFS |               |
+   |                         |           | (REQ)       |               |
+NS*| CB_WANTS_CANCELLED      | OPT       | FDELG,      | Section 20.10 |
+   |                         |           | DDELG, pNFS |               |
+   |                         |           | (REQ)       |               |
+   +-------------------------+-----------+-------------+---------------+
+
+Implementation notes:
+
+SSV:
+* The spec claims this is mandatory, but we don't actually know of any
+  implementations, so we're ignoring it for now.  The server returns
+  NFS4ERR_ENCR_ALG_UNSUPP on EXCHANGE_ID, which should be future-proof.
+
+GSS on the backchannel:
+* Again, theoretically required but not widely implemented (in
+  particular, the current Linux client doesn't request it).  We return
+  NFS4ERR_ENCR_ALG_UNSUPP on CREATE_SESSION.
+
+DELEGPURGE:
+* mandatory only for servers that support CLAIM_DELEGATE_PREV and/or
+  CLAIM_DELEG_PREV_FH (which allows clients to keep delegations that
+  persist across client reboots).  Thus we need not implement this for
+  now.
+
+EXCHANGE_ID:
+* implementation ids are ignored
+
+CREATE_SESSION:
+* backchannel attributes are ignored
+
+SEQUENCE:
+* no support for dynamic slot table renegotiation (optional)
+
+Nonstandard compound limitations:
+* No support for a sessions fore channel RPC compound that requires both a
+  ca_maxrequestsize request and a ca_maxresponsesize reply, so we may
+  fail to live up to the promise we made in CREATE_SESSION fore channel
+  negotiation.
+
+See also http://wiki.linux-nfs.org/wiki/index.php/Server_4.0_and_4.1_issues.
diff --git a/Documentation/filesystems/nfs/nfsd-admin-interfaces.txt b/Documentation/filesystems/nfs/nfsd-admin-interfaces.txt
new file mode 100644
index 000000000..56a96fb08
--- /dev/null
+++ b/Documentation/filesystems/nfs/nfsd-admin-interfaces.txt
@@ -0,0 +1,41 @@
+Administrative interfaces for nfsd
+^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^^
+
+Note that normally these interfaces are used only by the utilities in
+nfs-utils.
+
+nfsd is controlled mainly by pseudofiles under the "nfsd" filesystem,
+which is normally mounted at /proc/fs/nfsd/.
+
+The server is always started by the first write of a nonzero value to
+nfsd/threads.
+
+Before doing that, NFSD can be told which sockets to listen on by
+writing to nfsd/portlist; that write may be:
+
+	- an ascii-encoded file descriptor, which should refer to a
+	  bound (and listening, for tcp) socket, or
+	- "transportname port", where transportname is currently either
+	  "udp", "tcp", or "rdma".
+
+If nfsd is started without doing any of these, then it will create one
+udp and one tcp listener at port 2049 (see nfsd_init_socks).
+
+On startup, nfsd and lockd grace periods start.
+
+nfsd is shut down by a write of 0 to nfsd/threads.  All locks and state
+are thrown away at that point.
+
+Between startup and shutdown, the number of threads may be adjusted up
+or down by additional writes to nfsd/threads or by writes to
+nfsd/pool_threads.
+
+For more detail about files under nfsd/ and what they control, see
+fs/nfsd/nfsctl.c; most of them have detailed comments.
+
+Implementation notes
+^^^^^^^^^^^^^^^^^^^^
+
+Note that the rpc server requires the caller to serialize addition and
+removal of listening sockets, and startup and shutdown of the server.
+For nfsd this is done using nfsd_mutex.
diff --git a/Documentation/filesystems/nfs/nfsroot.txt b/Documentation/filesystems/nfs/nfsroot.txt
new file mode 100644
index 000000000..d2963123e
--- /dev/null
+++ b/Documentation/filesystems/nfs/nfsroot.txt
@@ -0,0 +1,355 @@
+Mounting the root filesystem via NFS (nfsroot)
+===============================================
+
+Written 1996 by Gero Kuhlmann <gero@gkminix.han.de>
+Updated 1997 by Martin Mares <mj@atrey.karlin.mff.cuni.cz>
+Updated 2006 by Nico Schottelius <nico-kernel-nfsroot@schottelius.org>
+Updated 2006 by Horms <horms@verge.net.au>
+Updated 2018 by Chris Novakovic <chris@chrisn.me.uk>
+
+
+
+In order to use a diskless system, such as an X-terminal or printer server
+for example, it is necessary for the root filesystem to be present on a
+non-disk device. This may be an initramfs (see Documentation/filesystems/
+ramfs-rootfs-initramfs.txt), a ramdisk (see Documentation/admin-guide/initrd.rst) or a
+filesystem mounted via NFS. The following text describes on how to use NFS
+for the root filesystem. For the rest of this text 'client' means the
+diskless system, and 'server' means the NFS server.
+
+
+
+
+1.) Enabling nfsroot capabilities
+    -----------------------------
+
+In order to use nfsroot, NFS client support needs to be selected as
+built-in during configuration. Once this has been selected, the nfsroot
+option will become available, which should also be selected.
+
+In the networking options, kernel level autoconfiguration can be selected,
+along with the types of autoconfiguration to support. Selecting all of
+DHCP, BOOTP and RARP is safe.
+
+
+
+
+2.) Kernel command line
+    -------------------
+
+When the kernel has been loaded by a boot loader (see below) it needs to be
+told what root fs device to use. And in the case of nfsroot, where to find
+both the server and the name of the directory on the server to mount as root.
+This can be established using the following kernel command line parameters:
+
+
+root=/dev/nfs
+
+  This is necessary to enable the pseudo-NFS-device. Note that it's not a
+  real device but just a synonym to tell the kernel to use NFS instead of
+  a real device.
+
+
+nfsroot=[<server-ip>:]<root-dir>[,<nfs-options>]
+
+  If the `nfsroot' parameter is NOT given on the command line,
+  the default "/tftpboot/%s" will be used.
+
+  <server-ip>	Specifies the IP address of the NFS server.
+		The default address is determined by the `ip' parameter
+		(see below). This parameter allows the use of different
+		servers for IP autoconfiguration and NFS.
+
+  <root-dir>	Name of the directory on the server to mount as root.
+		If there is a "%s" token in the string, it will be
+		replaced by the ASCII-representation of the client's
+		IP address.
+
+  <nfs-options>	Standard NFS options. All options are separated by commas.
+		The following defaults are used:
+			port		= as given by server portmap daemon
+			rsize		= 4096
+			wsize		= 4096
+			timeo		= 7
+			retrans		= 3
+			acregmin	= 3
+			acregmax	= 60
+			acdirmin	= 30
+			acdirmax	= 60
+			flags		= hard, nointr, noposix, cto, ac
+
+
+ip=<client-ip>:<server-ip>:<gw-ip>:<netmask>:<hostname>:<device>:<autoconf>:
+   <dns0-ip>:<dns1-ip>:<ntp0-ip>
+
+  This parameter tells the kernel how to configure IP addresses of devices
+  and also how to set up the IP routing table. It was originally called
+  `nfsaddrs', but now the boot-time IP configuration works independently of
+  NFS, so it was renamed to `ip' and the old name remained as an alias for
+  compatibility reasons.
+
+  If this parameter is missing from the kernel command line, all fields are
+  assumed to be empty, and the defaults mentioned below apply. In general
+  this means that the kernel tries to configure everything using
+  autoconfiguration.
+
+  The <autoconf> parameter can appear alone as the value to the `ip'
+  parameter (without all the ':' characters before).  If the value is
+  "ip=off" or "ip=none", no autoconfiguration will take place, otherwise
+  autoconfiguration will take place.  The most common way to use this
+  is "ip=dhcp".
+
+  <client-ip>	IP address of the client.
+
+  		Default:  Determined using autoconfiguration.
+
+  <server-ip>	IP address of the NFS server. If RARP is used to determine
+		the client address and this parameter is NOT empty only
+		replies from the specified server are accepted.
+
+		Only required for NFS root. That is autoconfiguration
+		will not be triggered if it is missing and NFS root is not
+		in operation.
+
+		Value is exported to /proc/net/pnp with the prefix "bootserver "
+		(see below).
+
+		Default: Determined using autoconfiguration.
+		         The address of the autoconfiguration server is used.
+
+  <gw-ip>	IP address of a gateway if the server is on a different subnet.
+
+		Default: Determined using autoconfiguration.
+
+  <netmask>	Netmask for local network interface. If unspecified
+		the netmask is derived from the client IP address assuming
+		classful addressing.
+
+		Default:  Determined using autoconfiguration.
+
+  <hostname>	Name of the client. If a '.' character is present, anything
+		before the first '.' is used as the client's hostname, and anything
+		after it is used as its NIS domain name. May be supplied by
+		autoconfiguration, but its absence will not trigger autoconfiguration.
+		If specified and DHCP is used, the user-provided hostname (and NIS
+		domain name, if present) will be carried in the DHCP request; this
+		may cause a DNS record to be created or updated for the client.
+
+  		Default: Client IP address is used in ASCII notation.
+
+  <device>	Name of network device to use.
+
+		Default: If the host only has one device, it is used.
+			 Otherwise the device is determined using
+			 autoconfiguration. This is done by sending
+			 autoconfiguration requests out of all devices,
+			 and using the device that received the first reply.
+
+  <autoconf>	Method to use for autoconfiguration. In the case of options
+                which specify multiple autoconfiguration protocols,
+		requests are sent using all protocols, and the first one
+		to reply is used.
+
+		Only autoconfiguration protocols that have been compiled
+		into the kernel will be used, regardless of the value of
+		this option.
+
+                  off or none: don't use autoconfiguration
+				(do static IP assignment instead)
+		  on or any:   use any protocol available in the kernel
+			       (default)
+		  dhcp:        use DHCP
+		  bootp:       use BOOTP
+		  rarp:        use RARP
+		  both:        use both BOOTP and RARP but not DHCP
+		               (old option kept for backwards compatibility)
+
+		if dhcp is used, the client identifier can be used by following
+		format "ip=dhcp,client-id-type,client-id-value"
+
+                Default: any
+
+  <dns0-ip>	IP address of primary nameserver.
+		Value is exported to /proc/net/pnp with the prefix "nameserver "
+		(see below).
+
+		Default: None if not using autoconfiguration; determined
+		automatically if using autoconfiguration.
+
+  <dns1-ip>	IP address of secondary nameserver.
+		See <dns0-ip>.
+
+  <ntp0-ip>	IP address of a Network Time Protocol (NTP) server.
+		Value is exported to /proc/net/ipconfig/ntp_servers, but is
+		otherwise unused (see below).
+
+		Default: None if not using autoconfiguration; determined
+		automatically if using autoconfiguration.
+
+  After configuration (whether manual or automatic) is complete, two files
+  are created in the following format; lines are omitted if their respective
+  value is empty following configuration:
+
+  - /proc/net/pnp:
+
+	#PROTO: <DHCP|BOOTP|RARP|MANUAL>	(depending on configuration method)
+	domain <dns-domain>			(if autoconfigured, the DNS domain)
+	nameserver <dns0-ip>			(primary name server IP)
+	nameserver <dns1-ip>			(secondary name server IP)
+	nameserver <dns2-ip>			(tertiary name server IP)
+	bootserver <server-ip>			(NFS server IP)
+
+  - /proc/net/ipconfig/ntp_servers:
+
+	<ntp0-ip>				(NTP server IP)
+	<ntp1-ip>				(NTP server IP)
+	<ntp2-ip>				(NTP server IP)
+
+  <dns-domain> and <dns2-ip> (in /proc/net/pnp) and <ntp1-ip> and <ntp2-ip>
+  (in /proc/net/ipconfig/ntp_servers) are requested during autoconfiguration;
+  they cannot be specified as part of the "ip=" kernel command line parameter.
+
+  Because the "domain" and "nameserver" options are recognised by DNS
+  resolvers, /etc/resolv.conf is often linked to /proc/net/pnp on systems
+  that use an NFS root filesystem.
+
+  Note that the kernel will not synchronise the system time with any NTP
+  servers it discovers; this is the responsibility of a user space process
+  (e.g. an initrd/initramfs script that passes the IP addresses listed in
+  /proc/net/ipconfig/ntp_servers to an NTP client before mounting the real
+  root filesystem if it is on NFS).
+
+
+nfsrootdebug
+
+  This parameter enables debugging messages to appear in the kernel
+  log at boot time so that administrators can verify that the correct
+  NFS mount options, server address, and root path are passed to the
+  NFS client.
+
+
+rdinit=<executable file>
+
+  To specify which file contains the program that starts system
+  initialization, administrators can use this command line parameter.
+  The default value of this parameter is "/init".  If the specified
+  file exists and the kernel can execute it, root filesystem related
+  kernel command line parameters, including `nfsroot=', are ignored.
+
+  A description of the process of mounting the root file system can be
+  found in:
+
+    Documentation/early-userspace/README
+
+
+
+
+3.) Boot Loader
+    ----------
+
+To get the kernel into memory different approaches can be used.
+They depend on various facilities being available:
+
+
+3.1)  Booting from a floppy using syslinux
+
+	When building kernels, an easy way to create a boot floppy that uses
+	syslinux is to use the zdisk or bzdisk make targets which use zimage
+      	and bzimage images respectively. Both targets accept the
+     	FDARGS parameter which can be used to set the kernel command line.
+
+	e.g.
+	   make bzdisk FDARGS="root=/dev/nfs"
+
+   	Note that the user running this command will need to have
+     	access to the floppy drive device, /dev/fd0
+
+     	For more information on syslinux, including how to create bootdisks
+     	for prebuilt kernels, see http://syslinux.zytor.com/
+
+	N.B: Previously it was possible to write a kernel directly to
+	     a floppy using dd, configure the boot device using rdev, and
+	     boot using the resulting floppy. Linux no longer supports this
+	     method of booting.
+
+3.2) Booting from a cdrom using isolinux
+
+     	When building kernels, an easy way to create a bootable cdrom that
+     	uses isolinux is to use the isoimage target which uses a bzimage
+     	image. Like zdisk and bzdisk, this target accepts the FDARGS
+     	parameter which can be used to set the kernel command line.
+
+	e.g.
+	  make isoimage FDARGS="root=/dev/nfs"
+
+     	The resulting iso image will be arch/<ARCH>/boot/image.iso
+     	This can be written to a cdrom using a variety of tools including
+     	cdrecord.
+
+	e.g.
+	  cdrecord dev=ATAPI:1,0,0 arch/x86/boot/image.iso
+
+     	For more information on isolinux, including how to create bootdisks
+     	for prebuilt kernels, see http://syslinux.zytor.com/
+
+3.2) Using LILO
+	When using LILO all the necessary command line parameters may be
+	specified using the 'append=' directive in the LILO configuration
+	file.
+
+	However, to use the 'root=' directive you also need to create
+	a dummy root device, which may be removed after LILO is run.
+
+	mknod /dev/boot255 c 0 255
+
+	For information on configuring LILO, please refer to its documentation.
+
+3.3) Using GRUB
+	When using GRUB, kernel parameter are simply appended after the kernel
+	specification: kernel <kernel> <parameters>
+
+3.4) Using loadlin
+	loadlin may be used to boot Linux from a DOS command prompt without
+	requiring a local hard disk to mount as root. This has not been
+	thoroughly tested by the authors of this document, but in general
+	it should be possible configure the kernel command line similarly
+	to the configuration of LILO.
+
+	Please refer to the loadlin documentation for further information.
+
+3.5) Using a boot ROM
+	This is probably the most elegant way of booting a diskless client.
+	With a boot ROM the kernel is loaded using the TFTP protocol. The
+	authors of this document are not aware of any no commercial boot
+	ROMs that support booting Linux over the network. However, there
+	are two free implementations of a boot ROM, netboot-nfs and
+	etherboot, both of which are available on sunsite.unc.edu, and both
+	of which contain everything you need to boot a diskless Linux client.
+
+3.6) Using pxelinux
+	Pxelinux may be used to boot linux using the PXE boot loader
+	which is present on many modern network cards.
+
+	When using pxelinux, the kernel image is specified using
+	"kernel <relative-path-below /tftpboot>". The nfsroot parameters
+	are passed to the kernel by adding them to the "append" line.
+	It is common to use serial console in conjunction with pxeliunx,
+	see Documentation/admin-guide/serial-console.rst for more information.
+
+	For more information on isolinux, including how to create bootdisks
+	for prebuilt kernels, see http://syslinux.zytor.com/
+
+
+
+
+4.) Credits
+    -------
+
+  The nfsroot code in the kernel and the RARP support have been written
+  by Gero Kuhlmann <gero@gkminix.han.de>.
+
+  The rest of the IP layer autoconfiguration code has been written
+  by Martin Mares <mj@atrey.karlin.mff.cuni.cz>.
+
+  In order to write the initial version of nfsroot I would like to thank
+  Jens-Uwe Mager <jum@anubis.han.de> for his help.
diff --git a/Documentation/filesystems/nfs/pnfs-block-server.txt b/Documentation/filesystems/nfs/pnfs-block-server.txt
new file mode 100644
index 000000000..2143673cf
--- /dev/null
+++ b/Documentation/filesystems/nfs/pnfs-block-server.txt
@@ -0,0 +1,37 @@
+pNFS block layout server user guide
+
+The Linux NFS server now supports the pNFS block layout extension.  In this
+case the NFS server acts as Metadata Server (MDS) for pNFS, which in addition
+to handling all the metadata access to the NFS export also hands out layouts
+to the clients to directly access the underlying block devices that are
+shared with the client.
+
+To use pNFS block layouts with with the Linux NFS server the exported file
+system needs to support the pNFS block layouts (currently just XFS), and the
+file system must sit on shared storage (typically iSCSI) that is accessible
+to the clients in addition to the MDS.  As of now the file system needs to
+sit directly on the exported volume, striping or concatenation of
+volumes on the MDS and clients is not supported yet.
+
+On the server, pNFS block volume support is automatically if the file system
+support it.  On the client make sure the kernel has the CONFIG_PNFS_BLOCK
+option enabled, the blkmapd daemon from nfs-utils is running, and the
+file system is mounted using the NFSv4.1 protocol version (mount -o vers=4.1).
+
+If the nfsd server needs to fence a non-responding client it calls
+/sbin/nfsd-recall-failed with the first argument set to the IP address of
+the client, and the second argument set to the device node without the /dev
+prefix for the file system to be fenced. Below is an example file that shows
+how to translate the device into a serial number from SCSI EVPD 0x80:
+
+cat > /sbin/nfsd-recall-failed << EOF
+#!/bin/sh
+
+CLIENT="$1"
+DEV="/dev/$2"
+EVPD=`sg_inq --page=0x80 ${DEV} | \
+	grep "Unit serial number:" | \
+	awk -F ': ' '{print $2}'`
+
+echo "fencing client ${CLIENT} serial ${EVPD}" >> /var/log/pnfsd-fence.log
+EOF
diff --git a/Documentation/filesystems/nfs/pnfs-scsi-server.txt b/Documentation/filesystems/nfs/pnfs-scsi-server.txt
new file mode 100644
index 000000000..5bef7268b
--- /dev/null
+++ b/Documentation/filesystems/nfs/pnfs-scsi-server.txt
@@ -0,0 +1,23 @@
+
+pNFS SCSI layout server user guide
+==================================
+
+This document describes support for pNFS SCSI layouts in the Linux NFS server.
+With pNFS SCSI layouts, the NFS server acts as Metadata Server (MDS) for pNFS,
+which in addition to handling all the metadata access to the NFS export,
+also hands out layouts to the clients so that they can directly access the
+underlying SCSI LUNs that are shared with the client.
+
+To use pNFS SCSI layouts with with the Linux NFS server, the exported file
+system needs to support the pNFS SCSI layouts (currently just XFS), and the
+file system must sit on a SCSI LUN that is accessible to the clients in
+addition to the MDS.  As of now the file system needs to sit directly on the
+exported LUN, striping or concatenation of LUNs on the MDS and clients
+is not supported yet.
+
+On a server built with CONFIG_NFSD_SCSI, the pNFS SCSI volume support is
+automatically enabled if the file system is exported using the "pnfs"
+option and the underlying SCSI device support persistent reservations.
+On the client make sure the kernel has the CONFIG_PNFS_BLOCK option
+enabled, and the file system is mounted using the NFSv4.1 protocol
+version (mount -o vers=4.1).
diff --git a/Documentation/filesystems/nfs/pnfs.txt b/Documentation/filesystems/nfs/pnfs.txt
new file mode 100644
index 000000000..80dc0bdc3
--- /dev/null
+++ b/Documentation/filesystems/nfs/pnfs.txt
@@ -0,0 +1,73 @@
+Reference counting in pnfs:
+==========================
+
+The are several inter-related caches.  We have layouts which can
+reference multiple devices, each of which can reference multiple data servers.
+Each data server can be referenced by multiple devices.  Each device
+can be referenced by multiple layouts.  To keep all of this straight,
+we need to reference count.
+
+
+struct pnfs_layout_hdr
+----------------------
+The on-the-wire command LAYOUTGET corresponds to struct
+pnfs_layout_segment, usually referred to by the variable name lseg.
+Each nfs_inode may hold a pointer to a cache of these layout
+segments in nfsi->layout, of type struct pnfs_layout_hdr.
+
+We reference the header for the inode pointing to it, across each
+outstanding RPC call that references it (LAYOUTGET, LAYOUTRETURN,
+LAYOUTCOMMIT), and for each lseg held within.
+
+Each header is also (when non-empty) put on a list associated with
+struct nfs_client (cl_layouts).  Being put on this list does not bump
+the reference count, as the layout is kept around by the lseg that
+keeps it in the list.
+
+deviceid_cache
+--------------
+lsegs reference device ids, which are resolved per nfs_client and
+layout driver type.  The device ids are held in a RCU cache (struct
+nfs4_deviceid_cache).  The cache itself is referenced across each
+mount.  The entries (struct nfs4_deviceid) themselves are held across
+the lifetime of each lseg referencing them.
+
+RCU is used because the deviceid is basically a write once, read many
+data structure.  The hlist size of 32 buckets needs better
+justification, but seems reasonable given that we can have multiple
+deviceid's per filesystem, and multiple filesystems per nfs_client.
+
+The hash code is copied from the nfsd code base.  A discussion of
+hashing and variations of this algorithm can be found at:
+http://groups.google.com/group/comp.lang.c/browse_thread/thread/9522965e2b8d3809
+
+data server cache
+-----------------
+file driver devices refer to data servers, which are kept in a module
+level cache.  Its reference is held over the lifetime of the deviceid
+pointing to it.
+
+lseg
+----
+lseg maintains an extra reference corresponding to the NFS_LSEG_VALID
+bit which holds it in the pnfs_layout_hdr's list.  When the final lseg
+is removed from the pnfs_layout_hdr's list, the NFS_LAYOUT_DESTROYED
+bit is set, preventing any new lsegs from being added.
+
+layout drivers
+--------------
+
+PNFS utilizes what is called layout drivers. The STD defines 4 basic
+layout types: "files", "objects", "blocks", and "flexfiles". For each
+of these types there is a layout-driver with a common function-vectors
+table which are called by the nfs-client pnfs-core to implement the
+different layout types.
+
+Files-layout-driver code is in: fs/nfs/filelayout/.. directory
+Blocks-layout-driver code is in: fs/nfs/blocklayout/.. directory
+Flexfiles-layout-driver code is in: fs/nfs/flexfilelayout/.. directory
+
+blocks-layout setup
+-------------------
+
+TODO: Document the setup needs of the blocks layout driver
diff --git a/Documentation/filesystems/nfs/rpc-cache.txt b/Documentation/filesystems/nfs/rpc-cache.txt
new file mode 100644
index 000000000..ebcaaee21
--- /dev/null
+++ b/Documentation/filesystems/nfs/rpc-cache.txt
@@ -0,0 +1,202 @@
+	This document gives a brief introduction to the caching
+mechanisms in the sunrpc layer that is used, in particular,
+for NFS authentication.
+
+CACHES
+======
+The caching replaces the old exports table and allows for
+a wide variety of values to be caches.
+
+There are a number of caches that are similar in structure though
+quite possibly very different in content and use.  There is a corpus
+of common code for managing these caches.
+
+Examples of caches that are likely to be needed are:
+  - mapping from IP address to client name
+  - mapping from client name and filesystem to export options
+  - mapping from UID to list of GIDs, to work around NFS's limitation
+    of 16 gids.
+  - mappings between local UID/GID and remote UID/GID for sites that
+    do not have uniform uid assignment
+  - mapping from network identify to public key for crypto authentication.
+
+The common code handles such things as:
+   - general cache lookup with correct locking
+   - supporting 'NEGATIVE' as well as positive entries
+   - allowing an EXPIRED time on cache items, and removing
+     items after they expire, and are no longer in-use.
+   - making requests to user-space to fill in cache entries
+   - allowing user-space to directly set entries in the cache
+   - delaying RPC requests that depend on as-yet incomplete
+     cache entries, and replaying those requests when the cache entry
+     is complete.
+   - clean out old entries as they expire.
+
+Creating a Cache
+----------------
+
+1/ A cache needs a datum to store.  This is in the form of a
+   structure definition that must contain a
+     struct cache_head
+   as an element, usually the first.
+   It will also contain a key and some content.
+   Each cache element is reference counted and contains
+   expiry and update times for use in cache management.
+2/ A cache needs a "cache_detail" structure that
+   describes the cache.  This stores the hash table, some
+   parameters for cache management, and some operations detailing how
+   to work with particular cache items.
+   The operations requires are:
+   	struct cache_head *alloc(void)
+		This simply allocates appropriate memory and returns
+   		a pointer to the cache_detail embedded within the
+		structure
+	void cache_put(struct kref *)
+		This is called when the last reference to an item is
+		dropped.  The pointer passed is to the 'ref' field
+		in the cache_head.  cache_put should release any
+		references create by 'cache_init' and, if CACHE_VALID
+		is set, any references created by cache_update.
+		It should then release the memory allocated by
+   		'alloc'.
+        int match(struct cache_head *orig, struct cache_head *new)
+		test if the keys in the two structures match.  Return
+		1 if they do, 0 if they don't.
+	void init(struct cache_head *orig, struct cache_head *new)
+		Set the 'key' fields in 'new' from 'orig'.  This may
+		include taking references to shared objects.
+	void update(struct cache_head *orig, struct cache_head *new)
+		Set the 'content' fileds in 'new' from 'orig'.
+	int cache_show(struct seq_file *m, struct cache_detail *cd,
+			struct cache_head *h)
+		Optional.  Used to provide a /proc file that lists the
+		contents of a cache.  This should show one item,
+   		usually on just one line.
+	int cache_request(struct cache_detail *cd, struct cache_head *h,
+   		char **bpp, int *blen)
+		Format a request to be send to user-space for an item
+   		to be instantiated.  *bpp is a buffer of size *blen.
+		bpp should be moved forward over the encoded message,
+		and  *blen should be reduced to show how much free
+		space remains.  Return 0 on success or <0 if not
+		enough room or other problem.
+	int cache_parse(struct cache_detail *cd, char *buf, int len)
+		A message from user space has arrived to fill out a
+		cache entry.  It is in 'buf' of length 'len'.
+		cache_parse should parse this, find the item in the
+		cache with sunrpc_cache_lookup, and update the item
+		with sunrpc_cache_update.
+
+
+3/ A cache needs to be registered using cache_register().  This
+   includes it on a list of caches that will be regularly
+   cleaned to discard old data.
+
+Using a cache
+-------------
+
+To find a value in a cache, call sunrpc_cache_lookup passing a pointer
+to the cache_head in a sample item with the 'key' fields filled in.
+This will be passed to ->match to identify the target entry.  If no
+entry is found, a new entry will be create, added to the cache, and
+marked as not containing valid data.
+
+The item returned is typically passed to cache_check which will check
+if the data is valid, and may initiate an up-call to get fresh data.
+cache_check will return -ENOENT in the entry is negative or if an up
+call is needed but not possible, -EAGAIN if an upcall is pending,
+or 0 if the data is valid;
+
+cache_check can be passed a "struct cache_req *".  This structure is
+typically embedded in the actual request and can be used to create a
+deferred copy of the request (struct cache_deferred_req).  This is
+done when the found cache item is not uptodate, but the is reason to
+believe that userspace might provide information soon.  When the cache
+item does become valid, the deferred copy of the request will be
+revisited (->revisit).  It is expected that this method will
+reschedule the request for processing.
+
+The value returned by sunrpc_cache_lookup can also be passed to
+sunrpc_cache_update to set the content for the item.  A second item is
+passed which should hold the content.  If the item found by _lookup
+has valid data, then it is discarded and a new item is created.  This
+saves any user of an item from worrying about content changing while
+it is being inspected.  If the item found by _lookup does not contain
+valid data, then the content is copied across and CACHE_VALID is set.
+
+Populating a cache
+------------------
+
+Each cache has a name, and when the cache is registered, a directory
+with that name is created in /proc/net/rpc
+
+This directory contains a file called 'channel' which is a channel
+for communicating between kernel and user for populating the cache.
+This directory may later contain other files of interacting
+with the cache.
+
+The 'channel' works a bit like a datagram socket. Each 'write' is
+passed as a whole to the cache for parsing and interpretation.
+Each cache can treat the write requests differently, but it is
+expected that a message written will contain:
+  - a key
+  - an expiry time
+  - a content.
+with the intention that an item in the cache with the give key
+should be create or updated to have the given content, and the
+expiry time should be set on that item.
+
+Reading from a channel is a bit more interesting.  When a cache
+lookup fails, or when it succeeds but finds an entry that may soon
+expire, a request is lodged for that cache item to be updated by
+user-space.  These requests appear in the channel file.
+
+Successive reads will return successive requests.
+If there are no more requests to return, read will return EOF, but a
+select or poll for read will block waiting for another request to be
+added.
+
+Thus a user-space helper is likely to:
+  open the channel.
+    select for readable
+    read a request
+    write a response
+  loop.
+
+If it dies and needs to be restarted, any requests that have not been
+answered will still appear in the file and will be read by the new
+instance of the helper.
+
+Each cache should define a "cache_parse" method which takes a message
+written from user-space and processes it.  It should return an error
+(which propagates back to the write syscall) or 0.
+
+Each cache should also define a "cache_request" method which
+takes a cache item and encodes a request into the buffer
+provided.
+
+Note: If a cache has no active readers on the channel, and has had not
+active readers for more than 60 seconds, further requests will not be
+added to the channel but instead all lookups that do not find a valid
+entry will fail.  This is partly for backward compatibility: The
+previous nfs exports table was deemed to be authoritative and a
+failed lookup meant a definite 'no'.
+
+request/response format
+-----------------------
+
+While each cache is free to use its own format for requests
+and responses over channel, the following is recommended as
+appropriate and support routines are available to help:
+Each request or response record should be printable ASCII
+with precisely one newline character which should be at the end.
+Fields within the record should be separated by spaces, normally one.
+If spaces, newlines, or nul characters are needed in a field they
+much be quoted.  two mechanisms are available:
+1/ If a field begins '\x' then it must contain an even number of
+   hex digits, and pairs of these digits provide the bytes in the
+   field.
+2/ otherwise a \ in the field must be followed by 3 octal digits
+   which give the code for a byte.  Other characters are treated
+   as them selves.  At the very least, space, newline, nul, and
+   '\' must be quoted in this way.
diff --git a/Documentation/filesystems/nfs/rpc-server-gss.txt b/Documentation/filesystems/nfs/rpc-server-gss.txt
new file mode 100644
index 000000000..310bbbaf9
--- /dev/null
+++ b/Documentation/filesystems/nfs/rpc-server-gss.txt
@@ -0,0 +1,91 @@
+
+rpcsec_gss support for kernel RPC servers
+=========================================
+
+This document gives references to the standards and protocols used to
+implement RPCGSS authentication in kernel RPC servers such as the NFS
+server and the NFS client's NFSv4.0 callback server.  (But note that
+NFSv4.1 and higher don't require the client to act as a server for the
+purposes of authentication.)
+
+RPCGSS is specified in a few IETF documents:
+ - RFC2203 v1: http://tools.ietf.org/rfc/rfc2203.txt
+ - RFC5403 v2: http://tools.ietf.org/rfc/rfc5403.txt
+and there is a 3rd version  being proposed:
+ - http://tools.ietf.org/id/draft-williams-rpcsecgssv3.txt
+   (At draft n. 02 at the time of writing)
+
+Background
+----------
+
+The RPCGSS Authentication method describes a way to perform GSSAPI
+Authentication for NFS.  Although GSSAPI is itself completely mechanism
+agnostic, in many cases only the KRB5 mechanism is supported by NFS
+implementations.
+
+The Linux kernel, at the moment, supports only the KRB5 mechanism, and
+depends on GSSAPI extensions that are KRB5 specific.
+
+GSSAPI is a complex library, and implementing it completely in kernel is
+unwarranted. However GSSAPI operations are fundementally separable in 2
+parts:
+- initial context establishment
+- integrity/privacy protection (signing and encrypting of individual
+  packets)
+
+The former is more complex and policy-independent, but less
+performance-sensitive.  The latter is simpler and needs to be very fast.
+
+Therefore, we perform per-packet integrity and privacy protection in the
+kernel, but leave the initial context establishment to userspace.  We
+need upcalls to request userspace to perform context establishment.
+
+NFS Server Legacy Upcall Mechanism
+----------------------------------
+
+The classic upcall mechanism uses a custom text based upcall mechanism
+to talk to a custom daemon called rpc.svcgssd that is provide by the
+nfs-utils package.
+
+This upcall mechanism has 2 limitations:
+
+A) It can handle tokens that are no bigger than 2KiB
+
+In some Kerberos deployment GSSAPI tokens can be quite big, up and
+beyond 64KiB in size due to various authorization extensions attacked to
+the Kerberos tickets, that needs to be sent through the GSS layer in
+order to perform context establishment.
+
+B) It does not properly handle creds where the user is member of more
+than a few thousand groups (the current hard limit in the kernel is 65K
+groups) due to limitation on the size of the buffer that can be send
+back to the kernel (4KiB).
+
+NFS Server New RPC Upcall Mechanism
+-----------------------------------
+
+The newer upcall mechanism uses RPC over a unix socket to a daemon
+called gss-proxy, implemented by a userspace program called Gssproxy.
+
+The gss_proxy RPC protocol is currently documented here:
+
+	https://fedorahosted.org/gss-proxy/wiki/ProtocolDocumentation
+
+This upcall mechanism uses the kernel rpc client and connects to the gssproxy
+userspace program over a regular unix socket. The gssproxy protocol does not
+suffer from the size limitations of the legacy protocol.
+
+Negotiating Upcall Mechanisms
+-----------------------------
+
+To provide backward compatibility, the kernel defaults to using the
+legacy mechanism.  To switch to the new mechanism, gss-proxy must bind
+to /var/run/gssproxy.sock and then write "1" to
+/proc/net/rpc/use-gss-proxy.  If gss-proxy dies, it must repeat both
+steps.
+
+Once the upcall mechanism is chosen, it cannot be changed.  To prevent
+locking into the legacy mechanisms, the above steps must be performed
+before starting nfsd.  Whoever starts nfsd can guarantee this by reading
+from /proc/net/rpc/use-gss-proxy and checking that it contains a
+"1"--the read will block until gss-proxy has done its write to the file.
diff --git a/Documentation/filesystems/nilfs2.txt b/Documentation/filesystems/nilfs2.txt
new file mode 100644
index 000000000..f2f3f8592
--- /dev/null
+++ b/Documentation/filesystems/nilfs2.txt
@@ -0,0 +1,276 @@
+NILFS2
+------
+
+NILFS2 is a log-structured file system (LFS) supporting continuous
+snapshotting.  In addition to versioning capability of the entire file
+system, users can even restore files mistakenly overwritten or
+destroyed just a few seconds ago.  Since NILFS2 can keep consistency
+like conventional LFS, it achieves quick recovery after system
+crashes.
+
+NILFS2 creates a number of checkpoints every few seconds or per
+synchronous write basis (unless there is no change).  Users can select
+significant versions among continuously created checkpoints, and can
+change them into snapshots which will be preserved until they are
+changed back to checkpoints.
+
+There is no limit on the number of snapshots until the volume gets
+full.  Each snapshot is mountable as a read-only file system
+concurrently with its writable mount, and this feature is convenient
+for online backup.
+
+The userland tools are included in nilfs-utils package, which is
+available from the following download page.  At least "mkfs.nilfs2",
+"mount.nilfs2", "umount.nilfs2", and "nilfs_cleanerd" (so called
+cleaner or garbage collector) are required.  Details on the tools are
+described in the man pages included in the package.
+
+Project web page:    https://nilfs.sourceforge.io/
+Download page:       https://nilfs.sourceforge.io/en/download.html
+List info:           http://vger.kernel.org/vger-lists.html#linux-nilfs
+
+Caveats
+=======
+
+Features which NILFS2 does not support yet:
+
+	- atime
+	- extended attributes
+	- POSIX ACLs
+	- quotas
+	- fsck
+	- defragmentation
+
+Mount options
+=============
+
+NILFS2 supports the following mount options:
+(*) == default
+
+barrier(*)		This enables/disables the use of write barriers.  This
+nobarrier		requires an IO stack which can support barriers, and
+			if nilfs gets an error on a barrier write, it will
+			disable again with a warning.
+errors=continue		Keep going on a filesystem error.
+errors=remount-ro(*)	Remount the filesystem read-only on an error.
+errors=panic		Panic and halt the machine if an error occurs.
+cp=n			Specify the checkpoint-number of the snapshot to be
+			mounted.  Checkpoints and snapshots are listed by lscp
+			user command.  Only the checkpoints marked as snapshot
+			are mountable with this option.  Snapshot is read-only,
+			so a read-only mount option must be specified together.
+order=relaxed(*)	Apply relaxed order semantics that allows modified data
+			blocks to be written to disk without making a
+			checkpoint if no metadata update is going.  This mode
+			is equivalent to the ordered data mode of the ext3
+			filesystem except for the updates on data blocks still
+			conserve atomicity.  This will improve synchronous
+			write performance for overwriting.
+order=strict		Apply strict in-order semantics that preserves sequence
+			of all file operations including overwriting of data
+			blocks.  That means, it is guaranteed that no
+			overtaking of events occurs in the recovered file
+			system after a crash.
+norecovery		Disable recovery of the filesystem on mount.
+			This disables every write access on the device for
+			read-only mounts or snapshots.  This option will fail
+			for r/w mounts on an unclean volume.
+discard			This enables/disables the use of discard/TRIM commands.
+nodiscard(*)		The discard/TRIM commands are sent to the underlying
+			block device when blocks are freed.  This is useful
+			for SSD devices and sparse/thinly-provisioned LUNs.
+
+Ioctls
+======
+
+There is some NILFS2 specific functionality which can be accessed by applications
+through the system call interfaces. The list of all NILFS2 specific ioctls are
+shown in the table below.
+
+Table of NILFS2 specific ioctls
+..............................................................................
+ Ioctl			        Description
+ NILFS_IOCTL_CHANGE_CPMODE      Change mode of given checkpoint between
+			        checkpoint and snapshot state. This ioctl is
+			        used in chcp and mkcp utilities.
+
+ NILFS_IOCTL_DELETE_CHECKPOINT  Remove checkpoint from NILFS2 file system.
+			        This ioctl is used in rmcp utility.
+
+ NILFS_IOCTL_GET_CPINFO         Return info about requested checkpoints. This
+			        ioctl is used in lscp utility and by
+			        nilfs_cleanerd daemon.
+
+ NILFS_IOCTL_GET_CPSTAT         Return checkpoints statistics. This ioctl is
+			        used by lscp, rmcp utilities and by
+			        nilfs_cleanerd daemon.
+
+ NILFS_IOCTL_GET_SUINFO         Return segment usage info about requested
+			        segments. This ioctl is used in lssu,
+			        nilfs_resize utilities and by nilfs_cleanerd
+			        daemon.
+
+ NILFS_IOCTL_SET_SUINFO         Modify segment usage info of requested
+				segments. This ioctl is used by
+				nilfs_cleanerd daemon to skip unnecessary
+				cleaning operation of segments and reduce
+				performance penalty or wear of flash device
+				due to redundant move of in-use blocks.
+
+ NILFS_IOCTL_GET_SUSTAT         Return segment usage statistics. This ioctl
+			        is used in lssu, nilfs_resize utilities and
+			        by nilfs_cleanerd daemon.
+
+ NILFS_IOCTL_GET_VINFO          Return information on virtual block addresses.
+			        This ioctl is used by nilfs_cleanerd daemon.
+
+ NILFS_IOCTL_GET_BDESCS         Return information about descriptors of disk
+			        block numbers. This ioctl is used by
+			        nilfs_cleanerd daemon.
+
+ NILFS_IOCTL_CLEAN_SEGMENTS     Do garbage collection operation in the
+			        environment of requested parameters from
+			        userspace. This ioctl is used by
+			        nilfs_cleanerd daemon.
+
+ NILFS_IOCTL_SYNC               Make a checkpoint. This ioctl is used in
+			        mkcp utility.
+
+ NILFS_IOCTL_RESIZE             Resize NILFS2 volume. This ioctl is used
+			        by nilfs_resize utility.
+
+ NILFS_IOCTL_SET_ALLOC_RANGE    Define lower limit of segments in bytes and
+			        upper limit of segments in bytes. This ioctl
+			        is used by nilfs_resize utility.
+
+NILFS2 usage
+============
+
+To use nilfs2 as a local file system, simply:
+
+ # mkfs -t nilfs2 /dev/block_device
+ # mount -t nilfs2 /dev/block_device /dir
+
+This will also invoke the cleaner through the mount helper program
+(mount.nilfs2).
+
+Checkpoints and snapshots are managed by the following commands.
+Their manpages are included in the nilfs-utils package above.
+
+  lscp     list checkpoints or snapshots.
+  mkcp     make a checkpoint or a snapshot.
+  chcp     change an existing checkpoint to a snapshot or vice versa.
+  rmcp     invalidate specified checkpoint(s).
+
+To mount a snapshot,
+
+ # mount -t nilfs2 -r -o cp=<cno> /dev/block_device /snap_dir
+
+where <cno> is the checkpoint number of the snapshot.
+
+To unmount the NILFS2 mount point or snapshot, simply:
+
+ # umount /dir
+
+Then, the cleaner daemon is automatically shut down by the umount
+helper program (umount.nilfs2).
+
+Disk format
+===========
+
+A nilfs2 volume is equally divided into a number of segments except
+for the super block (SB) and segment #0.  A segment is the container
+of logs.  Each log is composed of summary information blocks, payload
+blocks, and an optional super root block (SR):
+
+   ______________________________________________________
+  | |SB| | Segment | Segment | Segment | ... | Segment | |
+  |_|__|_|____0____|____1____|____2____|_____|____N____|_|
+  0 +1K +4K       +8M       +16M      +24M  +(8MB x N)
+       .             .            (Typical offsets for 4KB-block)
+    .                  .
+  .______________________.
+  | log | log |... | log |
+  |__1__|__2__|____|__m__|
+        .       .
+      .               .
+    .                       .
+  .______________________________.
+  | Summary | Payload blocks  |SR|
+  |_blocks__|_________________|__|
+
+The payload blocks are organized per file, and each file consists of
+data blocks and B-tree node blocks:
+
+    |<---       File-A        --->|<---       File-B        --->|
+   _______________________________________________________________
+    | Data blocks | B-tree blocks | Data blocks | B-tree blocks | ...
+   _|_____________|_______________|_____________|_______________|_
+
+
+Since only the modified blocks are written in the log, it may have
+files without data blocks or B-tree node blocks.
+
+The organization of the blocks is recorded in the summary information
+blocks, which contains a header structure (nilfs_segment_summary), per
+file structures (nilfs_finfo), and per block structures (nilfs_binfo):
+
+  _________________________________________________________________________
+ | Summary | finfo | binfo | ... | binfo | finfo | binfo | ... | binfo |...
+ |_blocks__|___A___|_(A,1)_|_____|(A,Na)_|___B___|_(B,1)_|_____|(B,Nb)_|___
+
+
+The logs include regular files, directory files, symbolic link files
+and several meta data files.  The mata data files are the files used
+to maintain file system meta data.  The current version of NILFS2 uses
+the following meta data files:
+
+ 1) Inode file (ifile)             -- Stores on-disk inodes
+ 2) Checkpoint file (cpfile)       -- Stores checkpoints
+ 3) Segment usage file (sufile)    -- Stores allocation state of segments
+ 4) Data address translation file  -- Maps virtual block numbers to usual
+    (DAT)                             block numbers.  This file serves to
+                                      make on-disk blocks relocatable.
+
+The following figure shows a typical organization of the logs:
+
+  _________________________________________________________________________
+ | Summary | regular file | file  | ... | ifile | cpfile | sufile | DAT |SR|
+ |_blocks__|_or_directory_|_______|_____|_______|________|________|_____|__|
+
+
+To stride over segment boundaries, this sequence of files may be split
+into multiple logs.  The sequence of logs that should be treated as
+logically one log, is delimited with flags marked in the segment
+summary.  The recovery code of nilfs2 looks this boundary information
+to ensure atomicity of updates.
+
+The super root block is inserted for every checkpoints.  It includes
+three special inodes, inodes for the DAT, cpfile, and sufile.  Inodes
+of regular files, directories, symlinks and other special files, are
+included in the ifile.  The inode of ifile itself is included in the
+corresponding checkpoint entry in the cpfile.  Thus, the hierarchy
+among NILFS2 files can be depicted as follows:
+
+  Super block (SB)
+       |
+       v
+  Super root block (the latest cno=xx)
+       |-- DAT
+       |-- sufile
+       `-- cpfile
+              |-- ifile (cno=c1)
+              |-- ifile (cno=c2) ---- file (ino=i1)
+              :        :          |-- file (ino=i2)
+              `-- ifile (cno=xx)  |-- file (ino=i3)
+                                  :        :
+                                  `-- file (ino=yy)
+                                    ( regular file, directory, or symlink )
+
+For detail on the format of each file, please see nilfs2_ondisk.h
+located at include/uapi/linux directory.
+
+There are no patents or other intellectual property that we protect
+with regard to the design of NILFS2.  It is allowed to replicate the
+design in hopes that other operating systems could share (mount, read,
+write, etc.) data stored in this format.
diff --git a/Documentation/filesystems/ntfs.txt b/Documentation/filesystems/ntfs.txt
new file mode 100644
index 000000000..553f10d03
--- /dev/null
+++ b/Documentation/filesystems/ntfs.txt
@@ -0,0 +1,451 @@
+The Linux NTFS filesystem driver
+================================
+
+
+Table of contents
+=================
+
+- Overview
+- Web site
+- Features
+- Supported mount options
+- Known bugs and (mis-)features
+- Using NTFS volume and stripe sets
+  - The Device-Mapper driver
+  - The Software RAID / MD driver
+  - Limitations when using the MD driver
+
+
+Overview
+========
+
+Linux-NTFS comes with a number of user-space programs known as ntfsprogs.
+These include mkntfs, a full-featured ntfs filesystem format utility,
+ntfsundelete used for recovering files that were unintentionally deleted
+from an NTFS volume and ntfsresize which is used to resize an NTFS partition.
+See the web site for more information.
+
+To mount an NTFS 1.2/3.x (Windows NT4/2000/XP/2003) volume, use the file
+system type 'ntfs'.  The driver currently supports read-only mode (with no
+fault-tolerance, encryption or journalling) and very limited, but safe, write
+support.
+
+For fault tolerance and raid support (i.e. volume and stripe sets), you can
+use the kernel's Software RAID / MD driver.  See section "Using Software RAID
+with NTFS" for details.
+
+
+Web site
+========
+
+There is plenty of additional information on the linux-ntfs web site
+at http://www.linux-ntfs.org/
+
+The web site has a lot of additional information, such as a comprehensive
+FAQ, documentation on the NTFS on-disk format, information on the Linux-NTFS
+userspace utilities, etc.
+
+
+Features
+========
+
+- This is a complete rewrite of the NTFS driver that used to be in the 2.4 and
+  earlier kernels.  This new driver implements NTFS read support and is
+  functionally equivalent to the old ntfs driver and it also implements limited
+  write support.  The biggest limitation at present is that files/directories
+  cannot be created or deleted.  See below for the list of write features that
+  are so far supported.  Another limitation is that writing to compressed files
+  is not implemented at all.  Also, neither read nor write access to encrypted
+  files is so far implemented.
+- The new driver has full support for sparse files on NTFS 3.x volumes which
+  the old driver isn't happy with.
+- The new driver supports execution of binaries due to mmap() now being
+  supported.
+- The new driver supports loopback mounting of files on NTFS which is used by
+  some Linux distributions to enable the user to run Linux from an NTFS
+  partition by creating a large file while in Windows and then loopback
+  mounting the file while in Linux and creating a Linux filesystem on it that
+  is used to install Linux on it.
+- A comparison of the two drivers using:
+	time find . -type f -exec md5sum "{}" \;
+  run three times in sequence with each driver (after a reboot) on a 1.4GiB
+  NTFS partition, showed the new driver to be 20% faster in total time elapsed
+  (from 9:43 minutes on average down to 7:53).  The time spent in user space
+  was unchanged but the time spent in the kernel was decreased by a factor of
+  2.5 (from 85 CPU seconds down to 33).
+- The driver does not support short file names in general.  For backwards
+  compatibility, we implement access to files using their short file names if
+  they exist.  The driver will not create short file names however, and a
+  rename will discard any existing short file name.
+- The new driver supports exporting of mounted NTFS volumes via NFS.
+- The new driver supports async io (aio).
+- The new driver supports fsync(2), fdatasync(2), and msync(2).
+- The new driver supports readv(2) and writev(2).
+- The new driver supports access time updates (including mtime and ctime).
+- The new driver supports truncate(2) and open(2) with O_TRUNC.  But at present
+  only very limited support for highly fragmented files, i.e. ones which have
+  their data attribute split across multiple extents, is included.  Another
+  limitation is that at present truncate(2) will never create sparse files,
+  since to mark a file sparse we need to modify the directory entry for the
+  file and we do not implement directory modifications yet.
+- The new driver supports write(2) which can both overwrite existing data and
+  extend the file size so that you can write beyond the existing data.  Also,
+  writing into sparse regions is supported and the holes are filled in with
+  clusters.  But at present only limited support for highly fragmented files,
+  i.e. ones which have their data attribute split across multiple extents, is
+  included.  Another limitation is that write(2) will never create sparse
+  files, since to mark a file sparse we need to modify the directory entry for
+  the file and we do not implement directory modifications yet.
+
+Supported mount options
+=======================
+
+In addition to the generic mount options described by the manual page for the
+mount command (man 8 mount, also see man 5 fstab), the NTFS driver supports the
+following mount options:
+
+iocharset=name		Deprecated option.  Still supported but please use
+			nls=name in the future.  See description for nls=name.
+
+nls=name		Character set to use when returning file names.
+			Unlike VFAT, NTFS suppresses names that contain
+			unconvertible characters.  Note that most character
+			sets contain insufficient characters to represent all
+			possible Unicode characters that can exist on NTFS.
+			To be sure you are not missing any files, you are
+			advised to use nls=utf8 which is capable of
+			representing all Unicode characters.
+
+utf8=<bool>		Option no longer supported.  Currently mapped to
+			nls=utf8 but please use nls=utf8 in the future and
+			make sure utf8 is compiled either as module or into
+			the kernel.  See description for nls=name.
+
+uid=
+gid=
+umask=			Provide default owner, group, and access mode mask.
+			These options work as documented in mount(8).  By
+			default, the files/directories are owned by root and
+			he/she has read and write permissions, as well as
+			browse permission for directories.  No one else has any
+			access permissions.  I.e. the mode on all files is by
+			default rw------- and for directories rwx------, a
+			consequence of the default fmask=0177 and dmask=0077.
+			Using a umask of zero will grant all permissions to
+			everyone, i.e. all files and directories will have mode
+			rwxrwxrwx.
+
+fmask=
+dmask=			Instead of specifying umask which applies both to
+			files and directories, fmask applies only to files and
+			dmask only to directories.
+
+sloppy=<BOOL>		If sloppy is specified, ignore unknown mount options.
+			Otherwise the default behaviour is to abort mount if
+			any unknown options are found.
+
+show_sys_files=<BOOL>	If show_sys_files is specified, show the system files
+			in directory listings.  Otherwise the default behaviour
+			is to hide the system files.
+			Note that even when show_sys_files is specified, "$MFT"
+			will not be visible due to bugs/mis-features in glibc.
+			Further, note that irrespective of show_sys_files, all
+			files are accessible by name, i.e. you can always do
+			"ls -l \$UpCase" for example to specifically show the
+			system file containing the Unicode upcase table.
+
+case_sensitive=<BOOL>	If case_sensitive is specified, treat all file names as
+			case sensitive and create file names in the POSIX
+			namespace.  Otherwise the default behaviour is to treat
+			file names as case insensitive and to create file names
+			in the WIN32/LONG name space.  Note, the Linux NTFS
+			driver will never create short file names and will
+			remove them on rename/delete of the corresponding long
+			file name.
+			Note that files remain accessible via their short file
+			name, if it exists.  If case_sensitive, you will need
+			to provide the correct case of the short file name.
+
+disable_sparse=<BOOL>	If disable_sparse is specified, creation of sparse
+			regions, i.e. holes, inside files is disabled for the
+			volume (for the duration of this mount only).  By
+			default, creation of sparse regions is enabled, which
+			is consistent with the behaviour of traditional Unix
+			filesystems.
+
+errors=opt		What to do when critical filesystem errors are found.
+			Following values can be used for "opt":
+			  continue: DEFAULT, try to clean-up as much as
+				    possible, e.g. marking a corrupt inode as
+				    bad so it is no longer accessed, and then
+				    continue.
+			  recover:  At present only supported is recovery of
+				    the boot sector from the backup copy.
+				    If read-only mount, the recovery is done
+				    in memory only and not written to disk.
+			Note that the options are additive, i.e. specifying:
+			   errors=continue,errors=recover
+			means the driver will attempt to recover and if that
+			fails it will clean-up as much as possible and
+			continue.
+
+mft_zone_multiplier=	Set the MFT zone multiplier for the volume (this
+			setting is not persistent across mounts and can be
+			changed from mount to mount but cannot be changed on
+			remount).  Values of 1 to 4 are allowed, 1 being the
+			default.  The MFT zone multiplier determines how much
+			space is reserved for the MFT on the volume.  If all
+			other space is used up, then the MFT zone will be
+			shrunk dynamically, so this has no impact on the
+			amount of free space.  However, it can have an impact
+			on performance by affecting fragmentation of the MFT.
+			In general use the default.  If you have a lot of small
+			files then use a higher value.  The values have the
+			following meaning:
+			      Value	     MFT zone size (% of volume size)
+				1		12.5%
+				2		25%
+				3		37.5%
+				4		50%
+			Note this option is irrelevant for read-only mounts.
+
+
+Known bugs and (mis-)features
+=============================
+
+- The link count on each directory inode entry is set to 1, due to Linux not
+  supporting directory hard links.  This may well confuse some user space
+  applications, since the directory names will have the same inode numbers.
+  This also speeds up ntfs_read_inode() immensely.  And we haven't found any
+  problems with this approach so far.  If you find a problem with this, please
+  let us know.
+
+
+Please send bug reports/comments/feedback/abuse to the Linux-NTFS development
+list at sourceforge: linux-ntfs-dev@lists.sourceforge.net
+
+
+Using NTFS volume and stripe sets
+=================================
+
+For support of volume and stripe sets, you can either use the kernel's
+Device-Mapper driver or the kernel's Software RAID / MD driver.  The former is
+the recommended one to use for linear raid.  But the latter is required for
+raid level 5.  For striping and mirroring, either driver should work fine.
+
+
+The Device-Mapper driver
+------------------------
+
+You will need to create a table of the components of the volume/stripe set and
+how they fit together and load this into the kernel using the dmsetup utility
+(see man 8 dmsetup).
+
+Linear volume sets, i.e. linear raid, has been tested and works fine.  Even
+though untested, there is no reason why stripe sets, i.e. raid level 0, and
+mirrors, i.e. raid level 1 should not work, too.  Stripes with parity, i.e.
+raid level 5, unfortunately cannot work yet because the current version of the
+Device-Mapper driver does not support raid level 5.  You may be able to use the
+Software RAID / MD driver for raid level 5, see the next section for details.
+
+To create the table describing your volume you will need to know each of its
+components and their sizes in sectors, i.e. multiples of 512-byte blocks.
+
+For NT4 fault tolerant volumes you can obtain the sizes using fdisk.  So for
+example if one of your partitions is /dev/hda2 you would do:
+
+$ fdisk -ul /dev/hda
+
+Disk /dev/hda: 81.9 GB, 81964302336 bytes
+255 heads, 63 sectors/track, 9964 cylinders, total 160086528 sectors
+Units = sectors of 1 * 512 = 512 bytes
+
+   Device Boot      Start         End      Blocks   Id  System
+   /dev/hda1   *          63     4209029     2104483+  83  Linux
+   /dev/hda2         4209030    37768814    16779892+  86  NTFS
+   /dev/hda3        37768815    46170809     4200997+  83  Linux
+
+And you would know that /dev/hda2 has a size of 37768814 - 4209030 + 1 =
+33559785 sectors.
+
+For Win2k and later dynamic disks, you can for example use the ldminfo utility
+which is part of the Linux LDM tools (the latest version at the time of
+writing is linux-ldm-0.0.8.tar.bz2).  You can download it from:
+	http://www.linux-ntfs.org/
+Simply extract the downloaded archive (tar xvjf linux-ldm-0.0.8.tar.bz2), go
+into it (cd linux-ldm-0.0.8) and change to the test directory (cd test).  You
+will find the precompiled (i386) ldminfo utility there.  NOTE: You will not be
+able to compile this yourself easily so use the binary version!
+
+Then you would use ldminfo in dump mode to obtain the necessary information:
+
+$ ./ldminfo --dump /dev/hda
+
+This would dump the LDM database found on /dev/hda which describes all of your
+dynamic disks and all the volumes on them.  At the bottom you will see the
+VOLUME DEFINITIONS section which is all you really need.  You may need to look
+further above to determine which of the disks in the volume definitions is
+which device in Linux.  Hint: Run ldminfo on each of your dynamic disks and
+look at the Disk Id close to the top of the output for each (the PRIVATE HEADER
+section).  You can then find these Disk Ids in the VBLK DATABASE section in the
+<Disk> components where you will get the LDM Name for the disk that is found in
+the VOLUME DEFINITIONS section.
+
+Note you will also need to enable the LDM driver in the Linux kernel.  If your
+distribution did not enable it, you will need to recompile the kernel with it
+enabled.  This will create the LDM partitions on each device at boot time.  You
+would then use those devices (for /dev/hda they would be /dev/hda1, 2, 3, etc)
+in the Device-Mapper table.
+
+You can also bypass using the LDM driver by using the main device (e.g.
+/dev/hda) and then using the offsets of the LDM partitions into this device as
+the "Start sector of device" when creating the table.  Once again ldminfo would
+give you the correct information to do this.
+
+Assuming you know all your devices and their sizes things are easy.
+
+For a linear raid the table would look like this (note all values are in
+512-byte sectors):
+
+--- cut here ---
+# Offset into	Size of this	Raid type	Device		Start sector
+# volume	device						of device
+0		1028161		linear		/dev/hda1	0
+1028161		3903762		linear		/dev/hdb2	0
+4931923		2103211		linear		/dev/hdc1	0
+--- cut here ---
+
+For a striped volume, i.e. raid level 0, you will need to know the chunk size
+you used when creating the volume.  Windows uses 64kiB as the default, so it
+will probably be this unless you changes the defaults when creating the array.
+
+For a raid level 0 the table would look like this (note all values are in
+512-byte sectors):
+
+--- cut here ---
+# Offset   Size	    Raid     Number   Chunk  1st        Start	2nd	  Start
+# into     of the   type     of	      size   Device	in	Device	  in
+# volume   volume	     stripes			device		  device
+0	   2056320  striped  2	      128    /dev/hda1	0	/dev/hdb1 0
+--- cut here ---
+
+If there are more than two devices, just add each of them to the end of the
+line.
+
+Finally, for a mirrored volume, i.e. raid level 1, the table would look like
+this (note all values are in 512-byte sectors):
+
+--- cut here ---
+# Ofs Size   Raid   Log  Number Region Should Number Source  Start Target Start
+# in  of the type   type of log size   sync?  of     Device  in    Device in
+# vol volume		 params		     mirrors	     Device	  Device
+0    2056320 mirror core 2	16     nosync 2	   /dev/hda1 0   /dev/hdb1 0
+--- cut here ---
+
+If you are mirroring to multiple devices you can specify further targets at the
+end of the line.
+
+Note the "Should sync?" parameter "nosync" means that the two mirrors are
+already in sync which will be the case on a clean shutdown of Windows.  If the
+mirrors are not clean, you can specify the "sync" option instead of "nosync"
+and the Device-Mapper driver will then copy the entirety of the "Source Device"
+to the "Target Device" or if you specified multiple target devices to all of
+them.
+
+Once you have your table, save it in a file somewhere (e.g. /etc/ntfsvolume1),
+and hand it over to dmsetup to work with, like so:
+
+$ dmsetup create myvolume1 /etc/ntfsvolume1
+
+You can obviously replace "myvolume1" with whatever name you like.
+
+If it all worked, you will now have the device /dev/device-mapper/myvolume1
+which you can then just use as an argument to the mount command as usual to
+mount the ntfs volume.  For example:
+
+$ mount -t ntfs -o ro /dev/device-mapper/myvolume1 /mnt/myvol1
+
+(You need to create the directory /mnt/myvol1 first and of course you can use
+anything you like instead of /mnt/myvol1 as long as it is an existing
+directory.)
+
+It is advisable to do the mount read-only to see if the volume has been setup
+correctly to avoid the possibility of causing damage to the data on the ntfs
+volume.
+
+
+The Software RAID / MD driver
+-----------------------------
+
+An alternative to using the Device-Mapper driver is to use the kernel's
+Software RAID / MD driver.  For which you need to set up your /etc/raidtab
+appropriately (see man 5 raidtab).
+
+Linear volume sets, i.e. linear raid, as well as stripe sets, i.e. raid level
+0, have been tested and work fine (though see section "Limitations when using
+the MD driver with NTFS volumes" especially if you want to use linear raid).
+Even though untested, there is no reason why mirrors, i.e. raid level 1, and
+stripes with parity, i.e. raid level 5, should not work, too.
+
+You have to use the "persistent-superblock 0" option for each raid-disk in the
+NTFS volume/stripe you are configuring in /etc/raidtab as the persistent
+superblock used by the MD driver would damage the NTFS volume.
+
+Windows by default uses a stripe chunk size of 64k, so you probably want the
+"chunk-size 64k" option for each raid-disk, too.
+
+For example, if you have a stripe set consisting of two partitions /dev/hda5
+and /dev/hdb1 your /etc/raidtab would look like this:
+
+raiddev /dev/md0
+	raid-level	0
+	nr-raid-disks	2
+	nr-spare-disks	0
+	persistent-superblock	0
+	chunk-size	64k
+	device		/dev/hda5
+	raid-disk	0
+	device		/dev/hdb1
+	raid-disk	1
+
+For linear raid, just change the raid-level above to "raid-level linear", for
+mirrors, change it to "raid-level 1", and for stripe sets with parity, change
+it to "raid-level 5".
+
+Note for stripe sets with parity you will also need to tell the MD driver
+which parity algorithm to use by specifying the option "parity-algorithm
+which", where you need to replace "which" with the name of the algorithm to
+use (see man 5 raidtab for available algorithms) and you will have to try the
+different available algorithms until you find one that works.  Make sure you
+are working read-only when playing with this as you may damage your data
+otherwise.  If you find which algorithm works please let us know (email the
+linux-ntfs developers list linux-ntfs-dev@lists.sourceforge.net or drop in on
+IRC in channel #ntfs on the irc.freenode.net network) so we can update this
+documentation.
+
+Once the raidtab is setup, run for example raid0run -a to start all devices or
+raid0run /dev/md0 to start a particular md device, in this case /dev/md0.
+
+Then just use the mount command as usual to mount the ntfs volume using for
+example:	mount -t ntfs -o ro /dev/md0 /mnt/myntfsvolume
+
+It is advisable to do the mount read-only to see if the md volume has been
+setup correctly to avoid the possibility of causing damage to the data on the
+ntfs volume.
+
+
+Limitations when using the Software RAID / MD driver
+-----------------------------------------------------
+
+Using the md driver will not work properly if any of your NTFS partitions have
+an odd number of sectors.  This is especially important for linear raid as all
+data after the first partition with an odd number of sectors will be offset by
+one or more sectors so if you mount such a partition with write support you
+will cause massive damage to the data on the volume which will only become
+apparent when you try to use the volume again under Windows.
+
+So when using linear raid, make sure that all your partitions have an even
+number of sectors BEFORE attempting to use it.  You have been warned!
+
+Even better is to simply use the Device-Mapper for linear raid and then you do
+not have this problem with odd numbers of sectors.
diff --git a/Documentation/filesystems/ocfs2-online-filecheck.txt b/Documentation/filesystems/ocfs2-online-filecheck.txt
new file mode 100644
index 000000000..139fab175
--- /dev/null
+++ b/Documentation/filesystems/ocfs2-online-filecheck.txt
@@ -0,0 +1,94 @@
+		    OCFS2 online file check
+		    -----------------------
+
+This document will describe OCFS2 online file check feature.
+
+Introduction
+============
+OCFS2 is often used in high-availability systems. However, OCFS2 usually
+converts the filesystem to read-only when encounters an error. This may not be
+necessary, since turning the filesystem read-only would affect other running
+processes as well, decreasing availability.
+Then, a mount option (errors=continue) is introduced, which would return the
+-EIO errno to the calling process and terminate further processing so that the
+filesystem is not corrupted further. The filesystem is not converted to
+read-only, and the problematic file's inode number is reported in the kernel
+log. The user can try to check/fix this file via online filecheck feature.
+
+Scope
+=====
+This effort is to check/fix small issues which may hinder day-to-day operations
+of a cluster filesystem by turning the filesystem read-only. The scope of
+checking/fixing is at the file level, initially for regular files and eventually
+to all files (including system files) of the filesystem.
+
+In case of directory to file links is incorrect, the directory inode is
+reported as erroneous.
+
+This feature is not suited for extravagant checks which involve dependency of
+other components of the filesystem, such as but not limited to, checking if the
+bits for file blocks in the allocation has been set. In case of such an error,
+the offline fsck should/would be recommended.
+
+Finally, such an operation/feature should not be automated lest the filesystem
+may end up with more damage than before the repair attempt. So, this has to
+be performed using user interaction and consent.
+
+User interface
+==============
+When there are errors in the OCFS2 filesystem, they are usually accompanied
+by the inode number which caused the error. This inode number would be the
+input to check/fix the file.
+
+There is a sysfs directory for each OCFS2 file system mounting:
+
+  /sys/fs/ocfs2/<devname>/filecheck
+
+Here, <devname> indicates the name of OCFS2 volume device which has been already
+mounted. The file above would accept inode numbers. This could be used to
+communicate with kernel space, tell which file(inode number) will be checked or
+fixed. Currently, three operations are supported, which includes checking
+inode, fixing inode and setting the size of result record history.
+
+1. If you want to know what error exactly happened to <inode> before fixing, do
+
+  # echo "<inode>" > /sys/fs/ocfs2/<devname>/filecheck/check
+  # cat /sys/fs/ocfs2/<devname>/filecheck/check
+
+The output is like this:
+  INO		DONE	ERROR
+39502		1	GENERATION
+
+<INO> lists the inode numbers.
+<DONE> indicates whether the operation has been finished.
+<ERROR> says what kind of errors was found. For the detailed error numbers,
+please refer to the file linux/fs/ocfs2/filecheck.h.
+
+2. If you determine to fix this inode, do
+
+  # echo "<inode>" > /sys/fs/ocfs2/<devname>/filecheck/fix
+  # cat /sys/fs/ocfs2/<devname>/filecheck/fix
+
+The output is like this:
+  INO		DONE	ERROR
+39502		1	SUCCESS
+
+This time, the <ERROR> column indicates whether this fix is successful or not.
+
+3. The record cache is used to store the history of check/fix results. It's
+default size is 10, and can be adjust between the range of 10 ~ 100. You can
+adjust the size like this:
+
+  # echo "<size>" > /sys/fs/ocfs2/<devname>/filecheck/set
+
+Fixing stuff
+============
+On receiving the inode, the filesystem would read the inode and the
+file metadata. In case of errors, the filesystem would fix the errors
+and report the problems it fixed in the kernel log. As a precautionary measure,
+the inode must first be checked for errors before performing a final fix.
+
+The inode and the result history will be maintained temporarily in a
+small linked list buffer which would contain the last (N) inodes
+fixed/checked, the detailed errors which were fixed/checked are printed in the
+kernel log.
diff --git a/Documentation/filesystems/ocfs2.txt b/Documentation/filesystems/ocfs2.txt
new file mode 100644
index 000000000..4c49e5410
--- /dev/null
+++ b/Documentation/filesystems/ocfs2.txt
@@ -0,0 +1,106 @@
+OCFS2 filesystem
+==================
+OCFS2 is a general purpose extent based shared disk cluster file
+system with many similarities to ext3. It supports 64 bit inode
+numbers, and has automatically extending metadata groups which may
+also make it attractive for non-clustered use.
+
+You'll want to install the ocfs2-tools package in order to at least
+get "mount.ocfs2" and "ocfs2_hb_ctl".
+
+Project web page:    http://ocfs2.wiki.kernel.org
+Tools git tree:      https://github.com/markfasheh/ocfs2-tools
+OCFS2 mailing lists: http://oss.oracle.com/projects/ocfs2/mailman/
+
+All code copyright 2005 Oracle except when otherwise noted.
+
+CREDITS:
+Lots of code taken from ext3 and other projects.
+
+Authors in alphabetical order:
+Joel Becker   <joel.becker@oracle.com>
+Zach Brown    <zach.brown@oracle.com>
+Mark Fasheh   <mfasheh@suse.com>
+Kurt Hackel   <kurt.hackel@oracle.com>
+Tao Ma        <tao.ma@oracle.com>
+Sunil Mushran <sunil.mushran@oracle.com>
+Manish Singh  <manish.singh@oracle.com>
+Tiger Yang    <tiger.yang@oracle.com>
+
+Caveats
+=======
+Features which OCFS2 does not support yet:
+	- Directory change notification (F_NOTIFY)
+	- Distributed Caching (F_SETLEASE/F_GETLEASE/break_lease)
+
+Mount options
+=============
+
+OCFS2 supports the following mount options:
+(*) == default
+
+barrier=1		This enables/disables barriers. barrier=0 disables it,
+			barrier=1 enables it.
+errors=remount-ro(*)	Remount the filesystem read-only on an error.
+errors=panic		Panic and halt the machine if an error occurs.
+intr		(*)	Allow signals to interrupt cluster operations.
+nointr			Do not allow signals to interrupt cluster
+			operations.
+noatime			Do not update access time.
+relatime(*)		Update atime if the previous atime is older than
+			mtime or ctime
+strictatime		Always update atime, but the minimum update interval
+			is specified by atime_quantum.
+atime_quantum=60(*)	OCFS2 will not update atime unless this number
+			of seconds has passed since the last update.
+			Set to zero to always update atime. This option need
+			work with strictatime.
+data=ordered	(*)	All data are forced directly out to the main file
+			system prior to its metadata being committed to the
+			journal.
+data=writeback		Data ordering is not preserved, data may be written
+			into the main file system after its metadata has been
+			committed to the journal.
+preferred_slot=0(*)	During mount, try to use this filesystem slot first. If
+			it is in use by another node, the first empty one found
+			will be chosen. Invalid values will be ignored.
+commit=nrsec	(*)	Ocfs2 can be told to sync all its data and metadata
+			every 'nrsec' seconds. The default value is 5 seconds.
+			This means that if you lose your power, you will lose
+			as much as the latest 5 seconds of work (your
+			filesystem will not be damaged though, thanks to the
+			journaling).  This default value (or any low value)
+			will hurt performance, but it's good for data-safety.
+			Setting it to 0 will have the same effect as leaving
+			it at the default (5 seconds).
+			Setting it to very large values will improve
+			performance.
+localalloc=8(*)		Allows custom localalloc size in MB. If the value is too
+			large, the fs will silently revert it to the default.
+localflocks		This disables cluster aware flock.
+inode64			Indicates that Ocfs2 is allowed to create inodes at
+			any location in the filesystem, including those which
+			will result in inode numbers occupying more than 32
+			bits of significance.
+user_xattr	(*)	Enables Extended User Attributes.
+nouser_xattr		Disables Extended User Attributes.
+acl			Enables POSIX Access Control Lists support.
+noacl		(*)	Disables POSIX Access Control Lists support.
+resv_level=2	(*)	Set how aggressive allocation reservations will be.
+			Valid values are between 0 (reservations off) to 8
+			(maximum space for reservations).
+dir_resv_level=	(*)	By default, directory reservations will scale with file
+			reservations - users should rarely need to change this
+			value. If allocation reservations are turned off, this
+			option will have no effect.
+coherency=full  (*)	Disallow concurrent O_DIRECT writes, cluster inode
+			lock will be taken to force other nodes drop cache,
+			therefore full cluster coherency is guaranteed even
+			for O_DIRECT writes.
+coherency=buffered	Allow concurrent O_DIRECT writes without EX lock among
+			nodes, which gains high performance at risk of getting
+			stale data on other nodes.
+journal_async_commit	Commit block can be written to disk without waiting
+			for descriptor blocks. If enabled older kernels cannot
+			mount the device. This will enable 'journal_checksum'
+			internally.
diff --git a/Documentation/filesystems/omfs.txt b/Documentation/filesystems/omfs.txt
new file mode 100644
index 000000000..1d0d41ff5
--- /dev/null
+++ b/Documentation/filesystems/omfs.txt
@@ -0,0 +1,106 @@
+Optimized MPEG Filesystem (OMFS)
+
+Overview
+========
+
+OMFS is a filesystem created by SonicBlue for use in the ReplayTV DVR
+and Rio Karma MP3 player.  The filesystem is extent-based, utilizing
+block sizes from 2k to 8k, with hash-based directories.  This
+filesystem driver may be used to read and write disks from these
+devices.
+
+Note, it is not recommended that this FS be used in place of a general
+filesystem for your own streaming media device.  Native Linux filesystems
+will likely perform better.
+
+More information is available at:
+
+    http://linux-karma.sf.net/
+
+Various utilities, including mkomfs and omfsck, are included with
+omfsprogs, available at:
+
+    http://bobcopeland.com/karma/
+
+Instructions are included in its README.
+
+Options
+=======
+
+OMFS supports the following mount-time options:
+
+    uid=n        - make all files owned by specified user
+    gid=n        - make all files owned by specified group
+    umask=xxx    - set permission umask to xxx
+    fmask=xxx    - set umask to xxx for files
+    dmask=xxx    - set umask to xxx for directories
+
+Disk format
+===========
+
+OMFS discriminates between "sysblocks" and normal data blocks.  The sysblock
+group consists of super block information, file metadata, directory structures,
+and extents.  Each sysblock has a header containing CRCs of the entire
+sysblock, and may be mirrored in successive blocks on the disk.  A sysblock may
+have a smaller size than a data block, but since they are both addressed by the
+same 64-bit block number, any remaining space in the smaller sysblock is
+unused.
+
+Sysblock header information:
+
+struct omfs_header {
+        __be64 h_self;                  /* FS block where this is located */
+        __be32 h_body_size;             /* size of useful data after header */
+        __be16 h_crc;                   /* crc-ccitt of body_size bytes */
+        char h_fill1[2];
+        u8 h_version;                   /* version, always 1 */
+        char h_type;                    /* OMFS_INODE_X */
+        u8 h_magic;                     /* OMFS_IMAGIC */
+        u8 h_check_xor;                 /* XOR of header bytes before this */
+        __be32 h_fill2;
+};
+
+Files and directories are both represented by omfs_inode:
+
+struct omfs_inode {
+        struct omfs_header i_head;      /* header */
+        __be64 i_parent;                /* parent containing this inode */
+        __be64 i_sibling;               /* next inode in hash bucket */
+        __be64 i_ctime;                 /* ctime, in milliseconds */
+        char i_fill1[35];
+        char i_type;                    /* OMFS_[DIR,FILE] */
+        __be32 i_fill2;
+        char i_fill3[64];
+        char i_name[OMFS_NAMELEN];      /* filename */
+        __be64 i_size;                  /* size of file, in bytes */
+};
+
+Directories in OMFS are implemented as a large hash table.  Filenames are
+hashed then prepended into the bucket list beginning at OMFS_DIR_START.
+Lookup requires hashing the filename, then seeking across i_sibling pointers
+until a match is found on i_name.  Empty buckets are represented by block
+pointers with all-1s (~0).
+
+A file is an omfs_inode structure followed by an extent table beginning at
+OMFS_EXTENT_START:
+
+struct omfs_extent_entry {
+        __be64 e_cluster;               /* start location of a set of blocks */
+        __be64 e_blocks;                /* number of blocks after e_cluster */
+};
+
+struct omfs_extent {
+        __be64 e_next;                  /* next extent table location */
+        __be32 e_extent_count;          /* total # extents in this table */
+        __be32 e_fill;
+        struct omfs_extent_entry e_entry;       /* start of extent entries */
+};
+
+Each extent holds the block offset followed by number of blocks allocated to
+the extent.  The final extent in each table is a terminator with e_cluster
+being ~0 and e_blocks being ones'-complement of the total number of blocks
+in the table.
+
+If this table overflows, a continuation inode is written and pointed to by
+e_next.  These have a header but lack the rest of the inode structure.
+
diff --git a/Documentation/filesystems/orangefs.txt b/Documentation/filesystems/orangefs.txt
new file mode 100644
index 000000000..f4ba94950
--- /dev/null
+++ b/Documentation/filesystems/orangefs.txt
@@ -0,0 +1,529 @@
+ORANGEFS
+========
+
+OrangeFS is an LGPL userspace scale-out parallel storage system. It is ideal
+for large storage problems faced by HPC, BigData, Streaming Video,
+Genomics, Bioinformatics.
+
+Orangefs, originally called PVFS, was first developed in 1993 by
+Walt Ligon and Eric Blumer as a parallel file system for Parallel
+Virtual Machine (PVM) as part of a NASA grant to study the I/O patterns
+of parallel programs.
+
+Orangefs features include:
+
+  * Distributes file data among multiple file servers
+  * Supports simultaneous access by multiple clients
+  * Stores file data and metadata on servers using local file system
+    and access methods
+  * Userspace implementation is easy to install and maintain
+  * Direct MPI support
+  * Stateless
+
+
+MAILING LIST ARCHIVES
+=====================
+
+http://lists.orangefs.org/pipermail/devel_lists.orangefs.org/
+
+
+MAILING LIST SUBMISSIONS
+========================
+
+devel@lists.orangefs.org
+
+
+DOCUMENTATION
+=============
+
+http://www.orangefs.org/documentation/
+
+
+USERSPACE FILESYSTEM SOURCE
+===========================
+
+http://www.orangefs.org/download
+
+Orangefs versions prior to 2.9.3 would not be compatible with the
+upstream version of the kernel client.
+
+
+RUNNING ORANGEFS ON A SINGLE SERVER
+===================================
+
+OrangeFS is usually run in large installations with multiple servers and
+clients, but a complete filesystem can be run on a single machine for
+development and testing.
+
+On Fedora, install orangefs and orangefs-server.
+
+dnf -y install orangefs orangefs-server
+
+There is an example server configuration file in
+/etc/orangefs/orangefs.conf.  Change localhost to your hostname if
+necessary.
+
+To generate a filesystem to run xfstests against, see below.
+
+There is an example client configuration file in /etc/pvfs2tab.  It is a
+single line.  Uncomment it and change the hostname if necessary.  This
+controls clients which use libpvfs2.  This does not control the
+pvfs2-client-core.
+
+Create the filesystem.
+
+pvfs2-server -f /etc/orangefs/orangefs.conf
+
+Start the server.
+
+systemctl start orangefs-server
+
+Test the server.
+
+pvfs2-ping -m /pvfsmnt
+
+Start the client.  The module must be compiled in or loaded before this
+point.
+
+systemctl start orangefs-client
+
+Mount the filesystem.
+
+mount -t pvfs2 tcp://localhost:3334/orangefs /pvfsmnt
+
+
+BUILDING ORANGEFS ON A SINGLE SERVER
+====================================
+
+Where OrangeFS cannot be installed from distribution packages, it may be
+built from source.
+
+You can omit --prefix if you don't care that things are sprinkled around
+in /usr/local.  As of version 2.9.6, OrangeFS uses Berkeley DB by
+default, we will probably be changing the default to LMDB soon.
+
+./configure --prefix=/opt/ofs --with-db-backend=lmdb
+
+make
+
+make install
+
+Create an orangefs config file.
+
+/opt/ofs/bin/pvfs2-genconfig /etc/pvfs2.conf
+
+Create an /etc/pvfs2tab file.
+
+echo tcp://localhost:3334/orangefs /pvfsmnt pvfs2 defaults,noauto 0 0 > \
+    /etc/pvfs2tab
+
+Create the mount point you specified in the tab file if needed.
+
+mkdir /pvfsmnt
+
+Bootstrap the server.
+
+/opt/ofs/sbin/pvfs2-server -f /etc/pvfs2.conf
+
+Start the server.
+
+/opt/osf/sbin/pvfs2-server /etc/pvfs2.conf
+
+Now the server should be running. Pvfs2-ls is a simple
+test to verify that the server is running.
+
+/opt/ofs/bin/pvfs2-ls /pvfsmnt
+
+If stuff seems to be working, load the kernel module and
+turn on the client core.
+
+/opt/ofs/sbin/pvfs2-client -p /opt/osf/sbin/pvfs2-client-core
+
+Mount your filesystem.
+
+mount -t pvfs2 tcp://localhost:3334/orangefs /pvfsmnt
+
+
+RUNNING XFSTESTS
+================
+
+It is useful to use a scratch filesystem with xfstests.  This can be
+done with only one server.
+
+Make a second copy of the FileSystem section in the server configuration
+file, which is /etc/orangefs/orangefs.conf.  Change the Name to scratch.
+Change the ID to something other than the ID of the first FileSystem
+section (2 is usually a good choice).
+
+Then there are two FileSystem sections: orangefs and scratch.
+
+This change should be made before creating the filesystem.
+
+pvfs2-server -f /etc/orangefs/orangefs.conf
+
+To run xfstests, create /etc/xfsqa.config.
+
+TEST_DIR=/orangefs
+TEST_DEV=tcp://localhost:3334/orangefs
+SCRATCH_MNT=/scratch
+SCRATCH_DEV=tcp://localhost:3334/scratch
+
+Then xfstests can be run
+
+./check -pvfs2
+
+
+OPTIONS
+=======
+
+The following mount options are accepted:
+
+  acl
+    Allow the use of Access Control Lists on files and directories.
+
+  intr
+    Some operations between the kernel client and the user space
+    filesystem can be interruptible, such as changes in debug levels
+    and the setting of tunable parameters.
+
+  local_lock
+    Enable posix locking from the perspective of "this" kernel. The
+    default file_operations lock action is to return ENOSYS. Posix
+    locking kicks in if the filesystem is mounted with -o local_lock.
+    Distributed locking is being worked on for the future.
+
+
+DEBUGGING
+=========
+
+If you want the debug (GOSSIP) statements in a particular
+source file (inode.c for example) go to syslog:
+
+  echo inode > /sys/kernel/debug/orangefs/kernel-debug
+
+No debugging (the default):
+
+  echo none > /sys/kernel/debug/orangefs/kernel-debug
+
+Debugging from several source files:
+
+  echo inode,dir > /sys/kernel/debug/orangefs/kernel-debug
+
+All debugging:
+
+  echo all > /sys/kernel/debug/orangefs/kernel-debug
+
+Get a list of all debugging keywords:
+
+  cat /sys/kernel/debug/orangefs/debug-help
+
+
+PROTOCOL BETWEEN KERNEL MODULE AND USERSPACE
+============================================
+
+Orangefs is a user space filesystem and an associated kernel module.
+We'll just refer to the user space part of Orangefs as "userspace"
+from here on out. Orangefs descends from PVFS, and userspace code
+still uses PVFS for function and variable names. Userspace typedefs
+many of the important structures. Function and variable names in
+the kernel module have been transitioned to "orangefs", and The Linux
+Coding Style avoids typedefs, so kernel module structures that
+correspond to userspace structures are not typedefed.
+
+The kernel module implements a pseudo device that userspace
+can read from and write to. Userspace can also manipulate the
+kernel module through the pseudo device with ioctl.
+
+THE BUFMAP:
+
+At startup userspace allocates two page-size-aligned (posix_memalign)
+mlocked memory buffers, one is used for IO and one is used for readdir
+operations. The IO buffer is 41943040 bytes and the readdir buffer is
+4194304 bytes. Each buffer contains logical chunks, or partitions, and
+a pointer to each buffer is added to its own PVFS_dev_map_desc structure
+which also describes its total size, as well as the size and number of
+the partitions.
+
+A pointer to the IO buffer's PVFS_dev_map_desc structure is sent to a
+mapping routine in the kernel module with an ioctl. The structure is
+copied from user space to kernel space with copy_from_user and is used
+to initialize the kernel module's "bufmap" (struct orangefs_bufmap), which
+then contains:
+
+  * refcnt - a reference counter
+  * desc_size - PVFS2_BUFMAP_DEFAULT_DESC_SIZE (4194304) - the IO buffer's
+    partition size, which represents the filesystem's block size and
+    is used for s_blocksize in super blocks.
+  * desc_count - PVFS2_BUFMAP_DEFAULT_DESC_COUNT (10) - the number of
+    partitions in the IO buffer.
+  * desc_shift - log2(desc_size), used for s_blocksize_bits in super blocks.
+  * total_size - the total size of the IO buffer.
+  * page_count - the number of 4096 byte pages in the IO buffer.
+  * page_array - a pointer to page_count * (sizeof(struct page*)) bytes
+    of kcalloced memory. This memory is used as an array of pointers
+    to each of the pages in the IO buffer through a call to get_user_pages.
+  * desc_array - a pointer to desc_count * (sizeof(struct orangefs_bufmap_desc))
+    bytes of kcalloced memory. This memory is further intialized:
+
+      user_desc is the kernel's copy of the IO buffer's ORANGEFS_dev_map_desc
+      structure. user_desc->ptr points to the IO buffer.
+
+      pages_per_desc = bufmap->desc_size / PAGE_SIZE
+      offset = 0
+
+        bufmap->desc_array[0].page_array = &bufmap->page_array[offset]
+        bufmap->desc_array[0].array_count = pages_per_desc = 1024
+        bufmap->desc_array[0].uaddr = (user_desc->ptr) + (0 * 1024 * 4096)
+        offset += 1024
+                           .
+                           .
+                           .
+        bufmap->desc_array[9].page_array = &bufmap->page_array[offset]
+        bufmap->desc_array[9].array_count = pages_per_desc = 1024
+        bufmap->desc_array[9].uaddr = (user_desc->ptr) +
+                                               (9 * 1024 * 4096)
+        offset += 1024
+
+  * buffer_index_array - a desc_count sized array of ints, used to
+    indicate which of the IO buffer's partitions are available to use.
+  * buffer_index_lock - a spinlock to protect buffer_index_array during update.
+  * readdir_index_array - a five (ORANGEFS_READDIR_DEFAULT_DESC_COUNT) element
+    int array used to indicate which of the readdir buffer's partitions are
+    available to use.
+  * readdir_index_lock - a spinlock to protect readdir_index_array during
+    update.
+
+OPERATIONS:
+
+The kernel module builds an "op" (struct orangefs_kernel_op_s) when it
+needs to communicate with userspace. Part of the op contains the "upcall"
+which expresses the request to userspace. Part of the op eventually
+contains the "downcall" which expresses the results of the request.
+
+The slab allocator is used to keep a cache of op structures handy.
+
+At init time the kernel module defines and initializes a request list
+and an in_progress hash table to keep track of all the ops that are
+in flight at any given time.
+
+Ops are stateful:
+
+ * unknown  - op was just initialized
+ * waiting  - op is on request_list (upward bound)
+ * inprogr  - op is in progress (waiting for downcall)
+ * serviced - op has matching downcall; ok
+ * purged   - op has to start a timer since client-core
+              exited uncleanly before servicing op
+ * given up - submitter has given up waiting for it
+
+When some arbitrary userspace program needs to perform a
+filesystem operation on Orangefs (readdir, I/O, create, whatever)
+an op structure is initialized and tagged with a distinguishing ID
+number. The upcall part of the op is filled out, and the op is
+passed to the "service_operation" function.
+
+Service_operation changes the op's state to "waiting", puts
+it on the request list, and signals the Orangefs file_operations.poll
+function through a wait queue. Userspace is polling the pseudo-device
+and thus becomes aware of the upcall request that needs to be read.
+
+When the Orangefs file_operations.read function is triggered, the
+request list is searched for an op that seems ready-to-process.
+The op is removed from the request list. The tag from the op and
+the filled-out upcall struct are copy_to_user'ed back to userspace.
+
+If any of these (and some additional protocol) copy_to_users fail,
+the op's state is set to "waiting" and the op is added back to
+the request list. Otherwise, the op's state is changed to "in progress",
+and the op is hashed on its tag and put onto the end of a list in the
+in_progress hash table at the index the tag hashed to.
+
+When userspace has assembled the response to the upcall, it
+writes the response, which includes the distinguishing tag, back to
+the pseudo device in a series of io_vecs. This triggers the Orangefs
+file_operations.write_iter function to find the op with the associated
+tag and remove it from the in_progress hash table. As long as the op's
+state is not "canceled" or "given up", its state is set to "serviced".
+The file_operations.write_iter function returns to the waiting vfs,
+and back to service_operation through wait_for_matching_downcall.
+
+Service operation returns to its caller with the op's downcall
+part (the response to the upcall) filled out.
+
+The "client-core" is the bridge between the kernel module and
+userspace. The client-core is a daemon. The client-core has an
+associated watchdog daemon. If the client-core is ever signaled
+to die, the watchdog daemon restarts the client-core. Even though
+the client-core is restarted "right away", there is a period of
+time during such an event that the client-core is dead. A dead client-core
+can't be triggered by the Orangefs file_operations.poll function.
+Ops that pass through service_operation during a "dead spell" can timeout
+on the wait queue and one attempt is made to recycle them. Obviously,
+if the client-core stays dead too long, the arbitrary userspace processes
+trying to use Orangefs will be negatively affected. Waiting ops
+that can't be serviced will be removed from the request list and
+have their states set to "given up". In-progress ops that can't
+be serviced will be removed from the in_progress hash table and
+have their states set to "given up".
+
+Readdir and I/O ops are atypical with respect to their payloads.
+
+  - readdir ops use the smaller of the two pre-allocated pre-partitioned
+    memory buffers. The readdir buffer is only available to userspace.
+    The kernel module obtains an index to a free partition before launching
+    a readdir op. Userspace deposits the results into the indexed partition
+    and then writes them to back to the pvfs device.
+
+  - io (read and write) ops use the larger of the two pre-allocated
+    pre-partitioned memory buffers. The IO buffer is accessible from
+    both userspace and the kernel module. The kernel module obtains an
+    index to a free partition before launching an io op. The kernel module
+    deposits write data into the indexed partition, to be consumed
+    directly by userspace. Userspace deposits the results of read
+    requests into the indexed partition, to be consumed directly
+    by the kernel module.
+
+Responses to kernel requests are all packaged in pvfs2_downcall_t
+structs. Besides a few other members, pvfs2_downcall_t contains a
+union of structs, each of which is associated with a particular
+response type.
+
+The several members outside of the union are:
+ - int32_t type - type of operation.
+ - int32_t status - return code for the operation.
+ - int64_t trailer_size - 0 unless readdir operation.
+ - char *trailer_buf - initialized to NULL, used during readdir operations.
+
+The appropriate member inside the union is filled out for any
+particular response.
+
+  PVFS2_VFS_OP_FILE_IO
+    fill a pvfs2_io_response_t
+
+  PVFS2_VFS_OP_LOOKUP
+    fill a PVFS_object_kref
+
+  PVFS2_VFS_OP_CREATE
+    fill a PVFS_object_kref
+
+  PVFS2_VFS_OP_SYMLINK
+    fill a PVFS_object_kref
+
+  PVFS2_VFS_OP_GETATTR
+    fill in a PVFS_sys_attr_s (tons of stuff the kernel doesn't need)
+    fill in a string with the link target when the object is a symlink.
+
+  PVFS2_VFS_OP_MKDIR
+    fill a PVFS_object_kref
+
+  PVFS2_VFS_OP_STATFS
+    fill a pvfs2_statfs_response_t with useless info <g>. It is hard for
+    us to know, in a timely fashion, these statistics about our
+    distributed network filesystem.
+
+  PVFS2_VFS_OP_FS_MOUNT
+    fill a pvfs2_fs_mount_response_t which is just like a PVFS_object_kref
+    except its members are in a different order and "__pad1" is replaced
+    with "id".
+
+  PVFS2_VFS_OP_GETXATTR
+    fill a pvfs2_getxattr_response_t
+
+  PVFS2_VFS_OP_LISTXATTR
+    fill a pvfs2_listxattr_response_t
+
+  PVFS2_VFS_OP_PARAM
+    fill a pvfs2_param_response_t
+
+  PVFS2_VFS_OP_PERF_COUNT
+    fill a pvfs2_perf_count_response_t
+
+  PVFS2_VFS_OP_FSKEY
+    file a pvfs2_fs_key_response_t
+
+  PVFS2_VFS_OP_READDIR
+    jamb everything needed to represent a pvfs2_readdir_response_t into
+    the readdir buffer descriptor specified in the upcall.
+
+Userspace uses writev() on /dev/pvfs2-req to pass responses to the requests
+made by the kernel side.
+
+A buffer_list containing:
+  - a pointer to the prepared response to the request from the
+    kernel (struct pvfs2_downcall_t).
+  - and also, in the case of a readdir request, a pointer to a
+    buffer containing descriptors for the objects in the target
+    directory.
+... is sent to the function (PINT_dev_write_list) which performs
+the writev.
+
+PINT_dev_write_list has a local iovec array: struct iovec io_array[10];
+
+The first four elements of io_array are initialized like this for all
+responses:
+
+  io_array[0].iov_base = address of local variable "proto_ver" (int32_t)
+  io_array[0].iov_len = sizeof(int32_t)
+
+  io_array[1].iov_base = address of global variable "pdev_magic" (int32_t)
+  io_array[1].iov_len = sizeof(int32_t)
+
+  io_array[2].iov_base = address of parameter "tag" (PVFS_id_gen_t)
+  io_array[2].iov_len = sizeof(int64_t)
+
+  io_array[3].iov_base = address of out_downcall member (pvfs2_downcall_t)
+                         of global variable vfs_request (vfs_request_t)
+  io_array[3].iov_len = sizeof(pvfs2_downcall_t)
+
+Readdir responses initialize the fifth element io_array like this:
+
+  io_array[4].iov_base = contents of member trailer_buf (char *)
+                         from out_downcall member of global variable
+                         vfs_request
+  io_array[4].iov_len = contents of member trailer_size (PVFS_size)
+                        from out_downcall member of global variable
+                        vfs_request
+
+Orangefs exploits the dcache in order to avoid sending redundant
+requests to userspace. We keep object inode attributes up-to-date with
+orangefs_inode_getattr. Orangefs_inode_getattr uses two arguments to
+help it decide whether or not to update an inode: "new" and "bypass".
+Orangefs keeps private data in an object's inode that includes a short
+timeout value, getattr_time, which allows any iteration of
+orangefs_inode_getattr to know how long it has been since the inode was
+updated. When the object is not new (new == 0) and the bypass flag is not
+set (bypass == 0) orangefs_inode_getattr returns without updating the inode
+if getattr_time has not timed out. Getattr_time is updated each time the
+inode is updated.
+
+Creation of a new object (file, dir, sym-link) includes the evaluation of
+its pathname, resulting in a negative directory entry for the object.
+A new inode is allocated and associated with the dentry, turning it from
+a negative dentry into a "productive full member of society". Orangefs
+obtains the new inode from Linux with new_inode() and associates
+the inode with the dentry by sending the pair back to Linux with
+d_instantiate().
+
+The evaluation of a pathname for an object resolves to its corresponding
+dentry. If there is no corresponding dentry, one is created for it in
+the dcache. Whenever a dentry is modified or verified Orangefs stores a
+short timeout value in the dentry's d_time, and the dentry will be trusted
+for that amount of time. Orangefs is a network filesystem, and objects
+can potentially change out-of-band with any particular Orangefs kernel module
+instance, so trusting a dentry is risky. The alternative to trusting
+dentries is to always obtain the needed information from userspace - at
+least a trip to the client-core, maybe to the servers. Obtaining information
+from a dentry is cheap, obtaining it from userspace is relatively expensive,
+hence the motivation to use the dentry when possible.
+
+The timeout values d_time and getattr_time are jiffy based, and the
+code is designed to avoid the jiffy-wrap problem:
+
+"In general, if the clock may have wrapped around more than once, there
+is no way to tell how much time has elapsed. However, if the times t1
+and t2 are known to be fairly close, we can reliably compute the
+difference in a way that takes into account the possibility that the
+clock may have wrapped between times."
+
+                      from course notes by instructor Andy Wang
+
diff --git a/Documentation/filesystems/overlayfs.txt b/Documentation/filesystems/overlayfs.txt
new file mode 100644
index 000000000..d7dc9c818
--- /dev/null
+++ b/Documentation/filesystems/overlayfs.txt
@@ -0,0 +1,483 @@
+Written by: Neil Brown
+Please see MAINTAINERS file for where to send questions.
+
+Overlay Filesystem
+==================
+
+This document describes a prototype for a new approach to providing
+overlay-filesystem functionality in Linux (sometimes referred to as
+union-filesystems).  An overlay-filesystem tries to present a
+filesystem which is the result over overlaying one filesystem on top
+of the other.
+
+
+Overlay objects
+---------------
+
+The overlay filesystem approach is 'hybrid', because the objects that
+appear in the filesystem do not always appear to belong to that filesystem.
+In many cases, an object accessed in the union will be indistinguishable
+from accessing the corresponding object from the original filesystem.
+This is most obvious from the 'st_dev' field returned by stat(2).
+
+While directories will report an st_dev from the overlay-filesystem,
+non-directory objects may report an st_dev from the lower filesystem or
+upper filesystem that is providing the object.  Similarly st_ino will
+only be unique when combined with st_dev, and both of these can change
+over the lifetime of a non-directory object.  Many applications and
+tools ignore these values and will not be affected.
+
+In the special case of all overlay layers on the same underlying
+filesystem, all objects will report an st_dev from the overlay
+filesystem and st_ino from the underlying filesystem.  This will
+make the overlay mount more compliant with filesystem scanners and
+overlay objects will be distinguishable from the corresponding
+objects in the original filesystem.
+
+On 64bit systems, even if all overlay layers are not on the same
+underlying filesystem, the same compliant behavior could be achieved
+with the "xino" feature.  The "xino" feature composes a unique object
+identifier from the real object st_ino and an underlying fsid index.
+If all underlying filesystems support NFS file handles and export file
+handles with 32bit inode number encoding (e.g. ext4), overlay filesystem
+will use the high inode number bits for fsid.  Even when the underlying
+filesystem uses 64bit inode numbers, users can still enable the "xino"
+feature with the "-o xino=on" overlay mount option.  That is useful for the
+case of underlying filesystems like xfs and tmpfs, which use 64bit inode
+numbers, but are very unlikely to use the high inode number bit.
+
+
+Upper and Lower
+---------------
+
+An overlay filesystem combines two filesystems - an 'upper' filesystem
+and a 'lower' filesystem.  When a name exists in both filesystems, the
+object in the 'upper' filesystem is visible while the object in the
+'lower' filesystem is either hidden or, in the case of directories,
+merged with the 'upper' object.
+
+It would be more correct to refer to an upper and lower 'directory
+tree' rather than 'filesystem' as it is quite possible for both
+directory trees to be in the same filesystem and there is no
+requirement that the root of a filesystem be given for either upper or
+lower.
+
+The lower filesystem can be any filesystem supported by Linux and does
+not need to be writable.  The lower filesystem can even be another
+overlayfs.  The upper filesystem will normally be writable and if it
+is it must support the creation of trusted.* extended attributes, and
+must provide valid d_type in readdir responses, so NFS is not suitable.
+
+A read-only overlay of two read-only filesystems may use any
+filesystem type.
+
+Directories
+-----------
+
+Overlaying mainly involves directories.  If a given name appears in both
+upper and lower filesystems and refers to a non-directory in either,
+then the lower object is hidden - the name refers only to the upper
+object.
+
+Where both upper and lower objects are directories, a merged directory
+is formed.
+
+At mount time, the two directories given as mount options "lowerdir" and
+"upperdir" are combined into a merged directory:
+
+  mount -t overlay overlay -olowerdir=/lower,upperdir=/upper,\
+  workdir=/work /merged
+
+The "workdir" needs to be an empty directory on the same filesystem
+as upperdir.
+
+Then whenever a lookup is requested in such a merged directory, the
+lookup is performed in each actual directory and the combined result
+is cached in the dentry belonging to the overlay filesystem.  If both
+actual lookups find directories, both are stored and a merged
+directory is created, otherwise only one is stored: the upper if it
+exists, else the lower.
+
+Only the lists of names from directories are merged.  Other content
+such as metadata and extended attributes are reported for the upper
+directory only.  These attributes of the lower directory are hidden.
+
+whiteouts and opaque directories
+--------------------------------
+
+In order to support rm and rmdir without changing the lower
+filesystem, an overlay filesystem needs to record in the upper filesystem
+that files have been removed.  This is done using whiteouts and opaque
+directories (non-directories are always opaque).
+
+A whiteout is created as a character device with 0/0 device number.
+When a whiteout is found in the upper level of a merged directory, any
+matching name in the lower level is ignored, and the whiteout itself
+is also hidden.
+
+A directory is made opaque by setting the xattr "trusted.overlay.opaque"
+to "y".  Where the upper filesystem contains an opaque directory, any
+directory in the lower filesystem with the same name is ignored.
+
+readdir
+-------
+
+When a 'readdir' request is made on a merged directory, the upper and
+lower directories are each read and the name lists merged in the
+obvious way (upper is read first, then lower - entries that already
+exist are not re-added).  This merged name list is cached in the
+'struct file' and so remains as long as the file is kept open.  If the
+directory is opened and read by two processes at the same time, they
+will each have separate caches.  A seekdir to the start of the
+directory (offset 0) followed by a readdir will cause the cache to be
+discarded and rebuilt.
+
+This means that changes to the merged directory do not appear while a
+directory is being read.  This is unlikely to be noticed by many
+programs.
+
+seek offsets are assigned sequentially when the directories are read.
+Thus if
+
+  - read part of a directory
+  - remember an offset, and close the directory
+  - re-open the directory some time later
+  - seek to the remembered offset
+
+there may be little correlation between the old and new locations in
+the list of filenames, particularly if anything has changed in the
+directory.
+
+Readdir on directories that are not merged is simply handled by the
+underlying directory (upper or lower).
+
+renaming directories
+--------------------
+
+When renaming a directory that is on the lower layer or merged (i.e. the
+directory was not created on the upper layer to start with) overlayfs can
+handle it in two different ways:
+
+1. return EXDEV error: this error is returned by rename(2) when trying to
+   move a file or directory across filesystem boundaries.  Hence
+   applications are usually prepared to hande this error (mv(1) for example
+   recursively copies the directory tree).  This is the default behavior.
+
+2. If the "redirect_dir" feature is enabled, then the directory will be
+   copied up (but not the contents).  Then the "trusted.overlay.redirect"
+   extended attribute is set to the path of the original location from the
+   root of the overlay.  Finally the directory is moved to the new
+   location.
+
+There are several ways to tune the "redirect_dir" feature.
+
+Kernel config options:
+
+- OVERLAY_FS_REDIRECT_DIR:
+    If this is enabled, then redirect_dir is turned on by  default.
+- OVERLAY_FS_REDIRECT_ALWAYS_FOLLOW:
+    If this is enabled, then redirects are always followed by default. Enabling
+    this results in a less secure configuration.  Enable this option only when
+    worried about backward compatibility with kernels that have the redirect_dir
+    feature and follow redirects even if turned off.
+
+Module options (can also be changed through /sys/module/overlay/parameters/*):
+
+- "redirect_dir=BOOL":
+    See OVERLAY_FS_REDIRECT_DIR kernel config option above.
+- "redirect_always_follow=BOOL":
+    See OVERLAY_FS_REDIRECT_ALWAYS_FOLLOW kernel config option above.
+- "redirect_max=NUM":
+    The maximum number of bytes in an absolute redirect (default is 256).
+
+Mount options:
+
+- "redirect_dir=on":
+    Redirects are enabled.
+- "redirect_dir=follow":
+    Redirects are not created, but followed.
+- "redirect_dir=off":
+    Redirects are not created and only followed if "redirect_always_follow"
+    feature is enabled in the kernel/module config.
+- "redirect_dir=nofollow":
+    Redirects are not created and not followed (equivalent to "redirect_dir=off"
+    if "redirect_always_follow" feature is not enabled).
+
+When the NFS export feature is enabled, every copied up directory is
+indexed by the file handle of the lower inode and a file handle of the
+upper directory is stored in a "trusted.overlay.upper" extended attribute
+on the index entry.  On lookup of a merged directory, if the upper
+directory does not match the file handle stores in the index, that is an
+indication that multiple upper directories may be redirected to the same
+lower directory.  In that case, lookup returns an error and warns about
+a possible inconsistency.
+
+Because lower layer redirects cannot be verified with the index, enabling
+NFS export support on an overlay filesystem with no upper layer requires
+turning off redirect follow (e.g. "redirect_dir=nofollow").
+
+
+Non-directories
+---------------
+
+Objects that are not directories (files, symlinks, device-special
+files etc.) are presented either from the upper or lower filesystem as
+appropriate.  When a file in the lower filesystem is accessed in a way
+the requires write-access, such as opening for write access, changing
+some metadata etc., the file is first copied from the lower filesystem
+to the upper filesystem (copy_up).  Note that creating a hard-link
+also requires copy_up, though of course creation of a symlink does
+not.
+
+The copy_up may turn out to be unnecessary, for example if the file is
+opened for read-write but the data is not modified.
+
+The copy_up process first makes sure that the containing directory
+exists in the upper filesystem - creating it and any parents as
+necessary.  It then creates the object with the same metadata (owner,
+mode, mtime, symlink-target etc.) and then if the object is a file, the
+data is copied from the lower to the upper filesystem.  Finally any
+extended attributes are copied up.
+
+Once the copy_up is complete, the overlay filesystem simply
+provides direct access to the newly created file in the upper
+filesystem - future operations on the file are barely noticed by the
+overlay filesystem (though an operation on the name of the file such as
+rename or unlink will of course be noticed and handled).
+
+
+Multiple lower layers
+---------------------
+
+Multiple lower layers can now be given using the the colon (":") as a
+separator character between the directory names.  For example:
+
+  mount -t overlay overlay -olowerdir=/lower1:/lower2:/lower3 /merged
+
+As the example shows, "upperdir=" and "workdir=" may be omitted.  In
+that case the overlay will be read-only.
+
+The specified lower directories will be stacked beginning from the
+rightmost one and going left.  In the above example lower1 will be the
+top, lower2 the middle and lower3 the bottom layer.
+
+
+Metadata only copy up
+--------------------
+
+When metadata only copy up feature is enabled, overlayfs will only copy
+up metadata (as opposed to whole file), when a metadata specific operation
+like chown/chmod is performed. Full file will be copied up later when
+file is opened for WRITE operation.
+
+In other words, this is delayed data copy up operation and data is copied
+up when there is a need to actually modify data.
+
+There are multiple ways to enable/disable this feature. A config option
+CONFIG_OVERLAY_FS_METACOPY can be set/unset to enable/disable this feature
+by default. Or one can enable/disable it at module load time with module
+parameter metacopy=on/off. Lastly, there is also a per mount option
+metacopy=on/off to enable/disable this feature per mount.
+
+Do not use metacopy=on with untrusted upper/lower directories. Otherwise
+it is possible that an attacker can create a handcrafted file with
+appropriate REDIRECT and METACOPY xattrs, and gain access to file on lower
+pointed by REDIRECT. This should not be possible on local system as setting
+"trusted." xattrs will require CAP_SYS_ADMIN. But it should be possible
+for untrusted layers like from a pen drive.
+
+Note: redirect_dir={off|nofollow|follow(*)} conflicts with metacopy=on, and
+results in an error.
+
+(*) redirect_dir=follow only conflicts with metacopy=on if upperdir=... is
+given.
+
+Sharing and copying layers
+--------------------------
+
+Lower layers may be shared among several overlay mounts and that is indeed
+a very common practice.  An overlay mount may use the same lower layer
+path as another overlay mount and it may use a lower layer path that is
+beneath or above the path of another overlay lower layer path.
+
+Using an upper layer path and/or a workdir path that are already used by
+another overlay mount is not allowed and may fail with EBUSY.  Using
+partially overlapping paths is not allowed and may fail with EBUSY.
+If files are accessed from two overlayfs mounts which share or overlap the
+upper layer and/or workdir path the behavior of the overlay is undefined,
+though it will not result in a crash or deadlock.
+
+Mounting an overlay using an upper layer path, where the upper layer path
+was previously used by another mounted overlay in combination with a
+different lower layer path, is allowed, unless the "inodes index" feature
+or "metadata only copy up" feature is enabled.
+
+With the "inodes index" feature, on the first time mount, an NFS file
+handle of the lower layer root directory, along with the UUID of the lower
+filesystem, are encoded and stored in the "trusted.overlay.origin" extended
+attribute on the upper layer root directory.  On subsequent mount attempts,
+the lower root directory file handle and lower filesystem UUID are compared
+to the stored origin in upper root directory.  On failure to verify the
+lower root origin, mount will fail with ESTALE.  An overlayfs mount with
+"inodes index" enabled will fail with EOPNOTSUPP if the lower filesystem
+does not support NFS export, lower filesystem does not have a valid UUID or
+if the upper filesystem does not support extended attributes.
+
+For "metadata only copy up" feature there is no verification mechanism at
+mount time. So if same upper is mounted with different set of lower, mount
+probably will succeed but expect the unexpected later on. So don't do it.
+
+It is quite a common practice to copy overlay layers to a different
+directory tree on the same or different underlying filesystem, and even
+to a different machine.  With the "inodes index" feature, trying to mount
+the copied layers will fail the verification of the lower root file handle.
+
+
+Non-standard behavior
+---------------------
+
+Overlayfs can now act as a POSIX compliant filesystem with the following
+features turned on:
+
+1) "redirect_dir"
+
+Enabled with the mount option or module option: "redirect_dir=on" or with
+the kernel config option CONFIG_OVERLAY_FS_REDIRECT_DIR=y.
+
+If this feature is disabled, then rename(2) on a lower or merged directory
+will fail with EXDEV ("Invalid cross-device link").
+
+2) "inode index"
+
+Enabled with the mount option or module option "index=on" or with the
+kernel config option CONFIG_OVERLAY_FS_INDEX=y.
+
+If this feature is disabled and a file with multiple hard links is copied
+up, then this will "break" the link.  Changes will not be propagated to
+other names referring to the same inode.
+
+3) "xino"
+
+Enabled with the mount option "xino=auto" or "xino=on", with the module
+option "xino_auto=on" or with the kernel config option
+CONFIG_OVERLAY_FS_XINO_AUTO=y.  Also implicitly enabled by using the same
+underlying filesystem for all layers making up the overlay.
+
+If this feature is disabled or the underlying filesystem doesn't have
+enough free bits in the inode number, then overlayfs will not be able to
+guarantee that the values of st_ino and st_dev returned by stat(2) and the
+value of d_ino returned by readdir(3) will act like on a normal filesystem.
+E.g. the value of st_dev may be different for two objects in the same
+overlay filesystem and the value of st_ino for directory objects may not be
+persistent and could change even while the overlay filesystem is mounted.
+
+
+Changes to underlying filesystems
+---------------------------------
+
+Offline changes, when the overlay is not mounted, are allowed to either
+the upper or the lower trees.
+
+Changes to the underlying filesystems while part of a mounted overlay
+filesystem are not allowed.  If the underlying filesystem is changed,
+the behavior of the overlay is undefined, though it will not result in
+a crash or deadlock.
+
+When the overlay NFS export feature is enabled, overlay filesystems
+behavior on offline changes of the underlying lower layer is different
+than the behavior when NFS export is disabled.
+
+On every copy_up, an NFS file handle of the lower inode, along with the
+UUID of the lower filesystem, are encoded and stored in an extended
+attribute "trusted.overlay.origin" on the upper inode.
+
+When the NFS export feature is enabled, a lookup of a merged directory,
+that found a lower directory at the lookup path or at the path pointed
+to by the "trusted.overlay.redirect" extended attribute, will verify
+that the found lower directory file handle and lower filesystem UUID
+match the origin file handle that was stored at copy_up time.  If a
+found lower directory does not match the stored origin, that directory
+will not be merged with the upper directory.
+
+
+
+NFS export
+----------
+
+When the underlying filesystems supports NFS export and the "nfs_export"
+feature is enabled, an overlay filesystem may be exported to NFS.
+
+With the "nfs_export" feature, on copy_up of any lower object, an index
+entry is created under the index directory.  The index entry name is the
+hexadecimal representation of the copy up origin file handle.  For a
+non-directory object, the index entry is a hard link to the upper inode.
+For a directory object, the index entry has an extended attribute
+"trusted.overlay.upper" with an encoded file handle of the upper
+directory inode.
+
+When encoding a file handle from an overlay filesystem object, the
+following rules apply:
+
+1. For a non-upper object, encode a lower file handle from lower inode
+2. For an indexed object, encode a lower file handle from copy_up origin
+3. For a pure-upper object and for an existing non-indexed upper object,
+   encode an upper file handle from upper inode
+
+The encoded overlay file handle includes:
+ - Header including path type information (e.g. lower/upper)
+ - UUID of the underlying filesystem
+ - Underlying filesystem encoding of underlying inode
+
+This encoding format is identical to the encoding format file handles that
+are stored in extended attribute "trusted.overlay.origin".
+
+When decoding an overlay file handle, the following steps are followed:
+
+1. Find underlying layer by UUID and path type information.
+2. Decode the underlying filesystem file handle to underlying dentry.
+3. For a lower file handle, lookup the handle in index directory by name.
+4. If a whiteout is found in index, return ESTALE. This represents an
+   overlay object that was deleted after its file handle was encoded.
+5. For a non-directory, instantiate a disconnected overlay dentry from the
+   decoded underlying dentry, the path type and index inode, if found.
+6. For a directory, use the connected underlying decoded dentry, path type
+   and index, to lookup a connected overlay dentry.
+
+Decoding a non-directory file handle may return a disconnected dentry.
+copy_up of that disconnected dentry will create an upper index entry with
+no upper alias.
+
+When overlay filesystem has multiple lower layers, a middle layer
+directory may have a "redirect" to lower directory.  Because middle layer
+"redirects" are not indexed, a lower file handle that was encoded from the
+"redirect" origin directory, cannot be used to find the middle or upper
+layer directory.  Similarly, a lower file handle that was encoded from a
+descendant of the "redirect" origin directory, cannot be used to
+reconstruct a connected overlay path.  To mitigate the cases of
+directories that cannot be decoded from a lower file handle, these
+directories are copied up on encode and encoded as an upper file handle.
+On an overlay filesystem with no upper layer this mitigation cannot be
+used NFS export in this setup requires turning off redirect follow (e.g.
+"redirect_dir=nofollow").
+
+The overlay filesystem does not support non-directory connectable file
+handles, so exporting with the 'subtree_check' exportfs configuration will
+cause failures to lookup files over NFS.
+
+When the NFS export feature is enabled, all directory index entries are
+verified on mount time to check that upper file handles are not stale.
+This verification may cause significant overhead in some cases.
+
+
+Testsuite
+---------
+
+There's a testsuite originally developed by David Howells and currently
+maintained by Amir Goldstein at:
+
+  https://github.com/amir73il/unionmount-testsuite.git
+
+Run as root:
+
+  # cd unionmount-testsuite
+  # ./run --ov --verify
diff --git a/Documentation/filesystems/path-lookup.md b/Documentation/filesystems/path-lookup.md
new file mode 100644
index 000000000..e2edd45c4
--- /dev/null
+++ b/Documentation/filesystems/path-lookup.md
@@ -0,0 +1,1297 @@
+<head>
+<style> p { max-width:50em} ol, ul {max-width: 40em}</style>
+</head>
+
+Pathname lookup in Linux.
+=========================
+
+This write-up is based on three articles published at lwn.net:
+
+- <https://lwn.net/Articles/649115/> Pathname lookup in Linux
+- <https://lwn.net/Articles/649729/> RCU-walk: faster pathname lookup in Linux
+- <https://lwn.net/Articles/650786/> A walk among the symlinks
+
+Written by Neil Brown with help from Al Viro and Jon Corbet.
+
+Introduction
+------------
+
+The most obvious aspect of pathname lookup, which very little
+exploration is needed to discover, is that it is complex.  There are
+many rules, special cases, and implementation alternatives that all
+combine to confuse the unwary reader.  Computer science has long been
+acquainted with such complexity and has tools to help manage it.  One
+tool that we will make extensive use of is "divide and conquer".  For
+the early parts of the analysis we will divide off symlinks - leaving
+them until the final part.  Well before we get to symlinks we have
+another major division based on the VFS's approach to locking which
+will allow us to review "REF-walk" and "RCU-walk" separately.  But we
+are getting ahead of ourselves.  There are some important low level
+distinctions we need to clarify first.
+
+There are two sorts of ...
+--------------------------
+
+[`openat()`]: http://man7.org/linux/man-pages/man2/openat.2.html
+
+Pathnames (sometimes "file names"), used to identify objects in the
+filesystem, will be familiar to most readers.  They contain two sorts
+of elements: "slashes" that are sequences of one or more "`/`"
+characters, and "components" that are sequences of one or more
+non-"`/`" characters.  These form two kinds of paths.  Those that
+start with slashes are "absolute" and start from the filesystem root.
+The others are "relative" and start from the current directory, or
+from some other location specified by a file descriptor given to a
+"xxx`at`" system call such as "[`openat()`]".
+
+[`execveat()`]: http://man7.org/linux/man-pages/man2/execveat.2.html
+
+It is tempting to describe the second kind as starting with a
+component, but that isn't always accurate: a pathname can lack both
+slashes and components, it can be empty, in other words.  This is
+generally forbidden in POSIX, but some of those "xxx`at`" system calls
+in Linux permit it when the `AT_EMPTY_PATH` flag is given.  For
+example, if you have an open file descriptor on an executable file you
+can execute it by calling [`execveat()`] passing the file descriptor,
+an empty path, and the `AT_EMPTY_PATH` flag.
+
+These paths can be divided into two sections: the final component and
+everything else.  The "everything else" is the easy bit.  In all cases
+it must identify a directory that already exists, otherwise an error
+such as `ENOENT` or `ENOTDIR` will be reported.
+
+The final component is not so simple.  Not only do different system
+calls interpret it quite differently (e.g. some create it, some do
+not), but it might not even exist: neither the empty pathname nor the
+pathname that is just slashes have a final component.  If it does
+exist, it could be "`.`" or "`..`" which are handled quite differently
+from other components.
+
+[POSIX]: http://pubs.opengroup.org/onlinepubs/9699919799/basedefs/V1_chap04.html#tag_04_12
+
+If a pathname ends with a slash, such as "`/tmp/foo/`" it might be
+tempting to consider that to have an empty final component.  In many
+ways that would lead to correct results, but not always.  In
+particular, `mkdir()` and `rmdir()` each create or remove a directory named
+by the final component, and they are required to work with pathnames
+ending in "`/`".  According to [POSIX]
+
+> A pathname that contains at least one non- &lt;slash> character and
+> that ends with one or more trailing &lt;slash> characters shall not
+> be resolved successfully unless the last pathname component before
+> the trailing <slash> characters names an existing directory or a
+> directory entry that is to be created for a directory immediately
+> after the pathname is resolved.
+
+The Linux pathname walking code (mostly in `fs/namei.c`) deals with
+all of these issues: breaking the path into components, handling the
+"everything else" quite separately from the final component, and
+checking that the trailing slash is not used where it isn't
+permitted.  It also addresses the important issue of concurrent
+access.
+
+While one process is looking up a pathname, another might be making
+changes that affect that lookup.  One fairly extreme case is that if
+"a/b" were renamed to "a/c/b" while another process were looking up
+"a/b/..", that process might successfully resolve on "a/c".
+Most races are much more subtle, and a big part of the task of
+pathname lookup is to prevent them from having damaging effects.  Many
+of the possible races are seen most clearly in the context of the
+"dcache" and an understanding of that is central to understanding
+pathname lookup.
+
+More than just a cache.
+-----------------------
+
+The "dcache" caches information about names in each filesystem to
+make them quickly available for lookup.  Each entry (known as a
+"dentry") contains three significant fields: a component name, a
+pointer to a parent dentry, and a pointer to the "inode" which
+contains further information about the object in that parent with
+the given name.  The inode pointer can be `NULL` indicating that the
+name doesn't exist in the parent.  While there can be linkage in the
+dentry of a directory to the dentries of the children, that linkage is
+not used for pathname lookup, and so will not be considered here.
+
+The dcache has a number of uses apart from accelerating lookup.  One
+that will be particularly relevant is that it is closely integrated
+with the mount table that records which filesystem is mounted where.
+What the mount table actually stores is which dentry is mounted on top
+of which other dentry.
+
+When considering the dcache, we have another of our "two types"
+distinctions: there are two types of filesystems.
+
+Some filesystems ensure that the information in the dcache is always
+completely accurate (though not necessarily complete).  This can allow
+the VFS to determine if a particular file does or doesn't exist
+without checking with the filesystem, and means that the VFS can
+protect the filesystem against certain races and other problems.
+These are typically "local" filesystems such as ext3, XFS, and Btrfs.
+
+Other filesystems don't provide that guarantee because they cannot.
+These are typically filesystems that are shared across a network,
+whether remote filesystems like NFS and 9P, or cluster filesystems
+like ocfs2 or cephfs.  These filesystems allow the VFS to revalidate
+cached information, and must provide their own protection against
+awkward races.  The VFS can detect these filesystems by the
+`DCACHE_OP_REVALIDATE` flag being set in the dentry.
+
+REF-walk: simple concurrency management with refcounts and spinlocks
+--------------------------------------------------------------------
+
+With all of those divisions carefully classified, we can now start
+looking at the actual process of walking along a path.  In particular
+we will start with the handling of the "everything else" part of a
+pathname, and focus on the "REF-walk" approach to concurrency
+management.  This code is found in the `link_path_walk()` function, if
+you ignore all the places that only run when "`LOOKUP_RCU`"
+(indicating the use of RCU-walk) is set.
+
+[Meet the Lockers]: https://lwn.net/Articles/453685/
+
+REF-walk is fairly heavy-handed with locks and reference counts.  Not
+as heavy-handed as in the old "big kernel lock" days, but certainly not
+afraid of taking a lock when one is needed.  It uses a variety of
+different concurrency controls.  A background understanding of the
+various primitives is assumed, or can be gleaned from elsewhere such
+as in [Meet the Lockers].
+
+The locking mechanisms used by REF-walk include:
+
+### dentry->d_lockref ###
+
+This uses the lockref primitive to provide both a spinlock and a
+reference count.  The special-sauce of this primitive is that the
+conceptual sequence "lock; inc_ref; unlock;" can often be performed
+with a single atomic memory operation.
+
+Holding a reference on a dentry ensures that the dentry won't suddenly
+be freed and used for something else, so the values in various fields
+will behave as expected.  It also protects the `->d_inode` reference
+to the inode to some extent.
+
+The association between a dentry and its inode is fairly permanent.
+For example, when a file is renamed, the dentry and inode move
+together to the new location.  When a file is created the dentry will
+initially be negative (i.e. `d_inode` is `NULL`), and will be assigned
+to the new inode as part of the act of creation.
+
+When a file is deleted, this can be reflected in the cache either by
+setting `d_inode` to `NULL`, or by removing it from the hash table
+(described shortly) used to look up the name in the parent directory.
+If the dentry is still in use the second option is used as it is
+perfectly legal to keep using an open file after it has been deleted
+and having the dentry around helps.  If the dentry is not otherwise in
+use (i.e. if the refcount in `d_lockref` is one), only then will
+`d_inode` be set to `NULL`.  Doing it this way is more efficient for a
+very common case.
+
+So as long as a counted reference is held to a dentry, a non-`NULL` `->d_inode`
+value will never be changed.
+
+### dentry->d_lock ###
+
+`d_lock` is a synonym for the spinlock that is part of `d_lockref` above.
+For our purposes, holding this lock protects against the dentry being
+renamed or unlinked.  In particular, its parent (`d_parent`), and its
+name (`d_name`) cannot be changed, and it cannot be removed from the
+dentry hash table.
+
+When looking for a name in a directory, REF-walk takes `d_lock` on
+each candidate dentry that it finds in the hash table and then checks
+that the parent and name are correct.  So it doesn't lock the parent
+while searching in the cache; it only locks children.
+
+When looking for the parent for a given name (to handle "`..`"),
+REF-walk can take `d_lock` to get a stable reference to `d_parent`,
+but it first tries a more lightweight approach.  As seen in
+`dget_parent()`, if a reference can be claimed on the parent, and if
+subsequently `d_parent` can be seen to have not changed, then there is
+no need to actually take the lock on the child.
+
+### rename_lock ###
+
+Looking up a given name in a given directory involves computing a hash
+from the two values (the name and the dentry of the directory),
+accessing that slot in a hash table, and searching the linked list
+that is found there.
+
+When a dentry is renamed, the name and the parent dentry can both
+change so the hash will almost certainly change too.  This would move the
+dentry to a different chain in the hash table.  If a filename search
+happened to be looking at a dentry that was moved in this way,
+it might end up continuing the search down the wrong chain,
+and so miss out on part of the correct chain.
+
+The name-lookup process (`d_lookup()`) does _not_ try to prevent this
+from happening, but only to detect when it happens.
+`rename_lock` is a seqlock that is updated whenever any dentry is
+renamed.  If `d_lookup` finds that a rename happened while it
+unsuccessfully scanned a chain in the hash table, it simply tries
+again.
+
+### inode->i_mutex ###
+
+`i_mutex` is a mutex that serializes all changes to a particular
+directory.  This ensures that, for example, an `unlink()` and a `rename()`
+cannot both happen at the same time.  It also keeps the directory
+stable while the filesystem is asked to look up a name that is not
+currently in the dcache.
+
+This has a complementary role to that of `d_lock`: `i_mutex` on a
+directory protects all of the names in that directory, while `d_lock`
+on a name protects just one name in a directory.  Most changes to the
+dcache hold `i_mutex` on the relevant directory inode and briefly take
+`d_lock` on one or more the dentries while the change happens.  One
+exception is when idle dentries are removed from the dcache due to
+memory pressure.  This uses `d_lock`, but `i_mutex` plays no role.
+
+The mutex affects pathname lookup in two distinct ways.  Firstly it
+serializes lookup of a name in a directory.  `walk_component()` uses
+`lookup_fast()` first which, in turn, checks to see if the name is in the cache,
+using only `d_lock` locking.  If the name isn't found, then `walk_component()`
+falls back to `lookup_slow()` which takes `i_mutex`, checks again that
+the name isn't in the cache, and then calls in to the filesystem to get a
+definitive answer.  A new dentry will be added to the cache regardless of
+the result.
+
+Secondly, when pathname lookup reaches the final component, it will
+sometimes need to take `i_mutex` before performing the last lookup so
+that the required exclusion can be achieved.  How path lookup chooses
+to take, or not take, `i_mutex` is one of the
+issues addressed in a subsequent section.
+
+### mnt->mnt_count ###
+
+`mnt_count` is a per-CPU reference counter on "`mount`" structures.
+Per-CPU here means that incrementing the count is cheap as it only
+uses CPU-local memory, but checking if the count is zero is expensive as
+it needs to check with every CPU.  Taking a `mnt_count` reference
+prevents the mount structure from disappearing as the result of regular
+unmount operations, but does not prevent a "lazy" unmount.  So holding
+`mnt_count` doesn't ensure that the mount remains in the namespace and,
+in particular, doesn't stabilize the link to the mounted-on dentry.  It
+does, however, ensure that the `mount` data structure remains coherent,
+and it provides a reference to the root dentry of the mounted
+filesystem.  So a reference through `->mnt_count` provides a stable
+reference to the mounted dentry, but not the mounted-on dentry.
+
+### mount_lock ###
+
+`mount_lock` is a global seqlock, a bit like `rename_lock`.  It can be used to
+check if any change has been made to any mount points.
+
+While walking down the tree (away from the root) this lock is used when
+crossing a mount point to check that the crossing was safe.  That is,
+the value in the seqlock is read, then the code finds the mount that
+is mounted on the current directory, if there is one, and increments
+the `mnt_count`.  Finally the value in `mount_lock` is checked against
+the old value.  If there is no change, then the crossing was safe.  If there
+was a change, the `mnt_count` is decremented and the whole process is
+retried.
+
+When walking up the tree (towards the root) by following a ".." link,
+a little more care is needed.  In this case the seqlock (which
+contains both a counter and a spinlock) is fully locked to prevent
+any changes to any mount points while stepping up.  This locking is
+needed to stabilize the link to the mounted-on dentry, which the
+refcount on the mount itself doesn't ensure.
+
+### RCU ###
+
+Finally the global (but extremely lightweight) RCU read lock is held
+from time to time to ensure certain data structures don't get freed
+unexpectedly.
+
+In particular it is held while scanning chains in the dcache hash
+table, and the mount point hash table.
+
+Bringing it together with `struct nameidata`
+--------------------------------------------
+
+[First edition Unix]: http://minnie.tuhs.org/cgi-bin/utree.pl?file=V1/u2.s
+
+Throughout the process of walking a path, the current status is stored
+in a `struct nameidata`, "namei" being the traditional name - dating
+all the way back to [First Edition Unix] - of the function that
+converts a "name" to an "inode".  `struct nameidata` contains (among
+other fields):
+
+### `struct path path` ###
+
+A `path` contains a `struct vfsmount` (which is
+embedded in a `struct mount`) and a `struct dentry`.  Together these
+record the current status of the walk.  They start out referring to the
+starting point (the current working directory, the root directory, or some other
+directory identified by a file descriptor), and are updated on each
+step.  A reference through `d_lockref` and `mnt_count` is always
+held.
+
+### `struct qstr last` ###
+
+This is a string together with a length (i.e. _not_ `nul` terminated)
+that is the "next" component in the pathname.
+
+### `int last_type` ###
+
+This is one of `LAST_NORM`, `LAST_ROOT`, `LAST_DOT`, `LAST_DOTDOT`, or
+`LAST_BIND`.  The `last` field is only valid if the type is
+`LAST_NORM`.  `LAST_BIND` is used when following a symlink and no
+components of the symlink have been processed yet.  Others should be
+fairly self-explanatory.
+
+### `struct path root` ###
+
+This is used to hold a reference to the effective root of the
+filesystem.  Often that reference won't be needed, so this field is
+only assigned the first time it is used, or when a non-standard root
+is requested.  Keeping a reference in the `nameidata` ensures that
+only one root is in effect for the entire path walk, even if it races
+with a `chroot()` system call.
+
+The root is needed when either of two conditions holds: (1) either the
+pathname or a symbolic link starts with a "'/'", or (2) a "`..`"
+component is being handled, since "`..`" from the root must always stay
+at the root.  The value used is usually the current root directory of
+the calling process.  An alternate root can be provided as when
+`sysctl()` calls `file_open_root()`, and when NFSv4 or Btrfs call
+`mount_subtree()`.  In each case a pathname is being looked up in a very
+specific part of the filesystem, and the lookup must not be allowed to
+escape that subtree.  It works a bit like a local `chroot()`.
+
+Ignoring the handling of symbolic links, we can now describe the
+"`link_path_walk()`" function, which handles the lookup of everything
+except the final component as:
+
+>  Given a path (`name`) and a nameidata structure (`nd`), check that the
+>  current directory has execute permission and then advance `name`
+>  over one component while updating `last_type` and `last`.  If that
+>  was the final component, then return, otherwise call
+>  `walk_component()` and repeat from the top.
+
+`walk_component()` is even easier.  If the component is `LAST_DOTS`,
+it calls `handle_dots()` which does the necessary locking as already
+described.  If it finds a `LAST_NORM` component it first calls
+"`lookup_fast()`" which only looks in the dcache, but will ask the
+filesystem to revalidate the result if it is that sort of filesystem.
+If that doesn't get a good result, it calls "`lookup_slow()`" which
+takes the `i_mutex`, rechecks the cache, and then asks the filesystem
+to find a definitive answer.  Each of these will call
+`follow_managed()` (as described below) to handle any mount points.
+
+In the absence of symbolic links, `walk_component()` creates a new
+`struct path` containing a counted reference to the new dentry and a
+reference to the new `vfsmount` which is only counted if it is
+different from the previous `vfsmount`.  It then calls
+`path_to_nameidata()` to install the new `struct path` in the
+`struct nameidata` and drop the unneeded references.
+
+This "hand-over-hand" sequencing of getting a reference to the new
+dentry before dropping the reference to the previous dentry may
+seem obvious, but is worth pointing out so that we will recognize its
+analogue in the "RCU-walk" version.
+
+Handling the final component.
+-----------------------------
+
+`link_path_walk()` only walks as far as setting `nd->last` and
+`nd->last_type` to refer to the final component of the path.  It does
+not call `walk_component()` that last time.  Handling that final
+component remains for the caller to sort out. Those callers are
+`path_lookupat()`, `path_parentat()`, `path_mountpoint()` and
+`path_openat()` each of which handles the differing requirements of
+different system calls.
+
+`path_parentat()` is clearly the simplest - it just wraps a little bit
+of housekeeping around `link_path_walk()` and returns the parent
+directory and final component to the caller.  The caller will be either
+aiming to create a name (via `filename_create()`) or remove or rename
+a name (in which case `user_path_parent()` is used).  They will use
+`i_mutex` to exclude other changes while they validate and then
+perform their operation.
+
+`path_lookupat()` is nearly as simple - it is used when an existing
+object is wanted such as by `stat()` or `chmod()`.  It essentially just
+calls `walk_component()` on the final component through a call to
+`lookup_last()`.  `path_lookupat()` returns just the final dentry.
+
+`path_mountpoint()` handles the special case of unmounting which must
+not try to revalidate the mounted filesystem.  It effectively
+contains, through a call to `mountpoint_last()`, an alternate
+implementation of `lookup_slow()` which skips that step.  This is
+important when unmounting a filesystem that is inaccessible, such as
+one provided by a dead NFS server.
+
+Finally `path_openat()` is used for the `open()` system call; it
+contains, in support functions starting with "`do_last()`", all the
+complexity needed to handle the different subtleties of O_CREAT (with
+or without O_EXCL), final "`/`" characters, and trailing symbolic
+links.  We will revisit this in the final part of this series, which
+focuses on those symbolic links.  "`do_last()`" will sometimes, but
+not always, take `i_mutex`, depending on what it finds.
+
+Each of these, or the functions which call them, need to be alert to
+the possibility that the final component is not `LAST_NORM`.  If the
+goal of the lookup is to create something, then any value for
+`last_type` other than `LAST_NORM` will result in an error.  For
+example if `path_parentat()` reports `LAST_DOTDOT`, then the caller
+won't try to create that name.  They also check for trailing slashes
+by testing `last.name[last.len]`.  If there is any character beyond
+the final component, it must be a trailing slash.
+
+Revalidation and automounts
+---------------------------
+
+Apart from symbolic links, there are only two parts of the "REF-walk"
+process not yet covered.  One is the handling of stale cache entries
+and the other is automounts.
+
+On filesystems that require it, the lookup routines will call the
+`->d_revalidate()` dentry method to ensure that the cached information
+is current.  This will often confirm validity or update a few details
+from a server.  In some cases it may find that there has been change
+further up the path and that something that was thought to be valid
+previously isn't really.  When this happens the lookup of the whole
+path is aborted and retried with the "`LOOKUP_REVAL`" flag set.  This
+forces revalidation to be more thorough.  We will see more details of
+this retry process in the next article.
+
+Automount points are locations in the filesystem where an attempt to
+lookup a name can trigger changes to how that lookup should be
+handled, in particular by mounting a filesystem there.  These are
+covered in greater detail in autofs.txt in the Linux documentation
+tree, but a few notes specifically related to path lookup are in order
+here.
+
+The Linux VFS has a concept of "managed" dentries which is reflected
+in function names such as "`follow_managed()`".  There are three
+potentially interesting things about these dentries corresponding
+to three different flags that might be set in `dentry->d_flags`:
+
+### `DCACHE_MANAGE_TRANSIT` ###
+
+If this flag has been set, then the filesystem has requested that the
+`d_manage()` dentry operation be called before handling any possible
+mount point.  This can perform two particular services:
+
+It can block to avoid races.  If an automount point is being
+unmounted, the `d_manage()` function will usually wait for that
+process to complete before letting the new lookup proceed and possibly
+trigger a new automount.
+
+It can selectively allow only some processes to transit through a
+mount point.  When a server process is managing automounts, it may
+need to access a directory without triggering normal automount
+processing.  That server process can identify itself to the `autofs`
+filesystem, which will then give it a special pass through
+`d_manage()` by returning `-EISDIR`.
+
+### `DCACHE_MOUNTED` ###
+
+This flag is set on every dentry that is mounted on.  As Linux
+supports multiple filesystem namespaces, it is possible that the
+dentry may not be mounted on in *this* namespace, just in some
+other.  So this flag is seen as a hint, not a promise.
+
+If this flag is set, and `d_manage()` didn't return `-EISDIR`,
+`lookup_mnt()` is called to examine the mount hash table (honoring the
+`mount_lock` described earlier) and possibly return a new `vfsmount`
+and a new `dentry` (both with counted references).
+
+### `DCACHE_NEED_AUTOMOUNT` ###
+
+If `d_manage()` allowed us to get this far, and `lookup_mnt()` didn't
+find a mount point, then this flag causes the `d_automount()` dentry
+operation to be called.
+
+The `d_automount()` operation can be arbitrarily complex and may
+communicate with server processes etc. but it should ultimately either
+report that there was an error, that there was nothing to mount, or
+should provide an updated `struct path` with new `dentry` and `vfsmount`.
+
+In the latter case, `finish_automount()` will be called to safely
+install the new mount point into the mount table.
+
+There is no new locking of import here and it is important that no
+locks (only counted references) are held over this processing due to
+the very real possibility of extended delays.
+This will become more important next time when we examine RCU-walk
+which is particularly sensitive to delays.
+
+RCU-walk - faster pathname lookup in Linux
+==========================================
+
+RCU-walk is another algorithm for performing pathname lookup in Linux.
+It is in many ways similar to REF-walk and the two share quite a bit
+of code.  The significant difference in RCU-walk is how it allows for
+the possibility of concurrent access.
+
+We noted that REF-walk is complex because there are numerous details
+and special cases.  RCU-walk reduces this complexity by simply
+refusing to handle a number of cases -- it instead falls back to
+REF-walk.  The difficulty with RCU-walk comes from a different
+direction: unfamiliarity.  The locking rules when depending on RCU are
+quite different from traditional locking, so we will spend a little extra
+time when we come to those.
+
+Clear demarcation of roles
+--------------------------
+
+The easiest way to manage concurrency is to forcibly stop any other
+thread from changing the data structures that a given thread is
+looking at.  In cases where no other thread would even think of
+changing the data and lots of different threads want to read at the
+same time, this can be very costly.  Even when using locks that permit
+multiple concurrent readers, the simple act of updating the count of
+the number of current readers can impose an unwanted cost.  So the
+goal when reading a shared data structure that no other process is
+changing is to avoid writing anything to memory at all.  Take no
+locks, increment no counts, leave no footprints.
+
+The REF-walk mechanism already described certainly doesn't follow this
+principle, but then it is really designed to work when there may well
+be other threads modifying the data.  RCU-walk, in contrast, is
+designed for the common situation where there are lots of frequent
+readers and only occasional writers.  This may not be common in all
+parts of the filesystem tree, but in many parts it will be.  For the
+other parts it is important that RCU-walk can quickly fall back to
+using REF-walk.
+
+Pathname lookup always starts in RCU-walk mode but only remains there
+as long as what it is looking for is in the cache and is stable.  It
+dances lightly down the cached filesystem image, leaving no footprints
+and carefully watching where it is, to be sure it doesn't trip.  If it
+notices that something has changed or is changing, or if something
+isn't in the cache, then it tries to stop gracefully and switch to
+REF-walk.
+
+This stopping requires getting a counted reference on the current
+`vfsmount` and `dentry`, and ensuring that these are still valid -
+that a path walk with REF-walk would have found the same entries.
+This is an invariant that RCU-walk must guarantee.  It can only make
+decisions, such as selecting the next step, that are decisions which
+REF-walk could also have made if it were walking down the tree at the
+same time.  If the graceful stop succeeds, the rest of the path is
+processed with the reliable, if slightly sluggish, REF-walk.  If
+RCU-walk finds it cannot stop gracefully, it simply gives up and
+restarts from the top with REF-walk.
+
+This pattern of "try RCU-walk, if that fails try REF-walk" can be
+clearly seen in functions like `filename_lookup()`,
+`filename_parentat()`, `filename_mountpoint()`,
+`do_filp_open()`, and `do_file_open_root()`.  These five
+correspond roughly to the four `path_`* functions we met earlier,
+each of which calls `link_path_walk()`.  The `path_*` functions are
+called using different mode flags until a mode is found which works.
+They are first called with `LOOKUP_RCU` set to request "RCU-walk".  If
+that fails with the error `ECHILD` they are called again with no
+special flag to request "REF-walk".  If either of those report the
+error `ESTALE` a final attempt is made with `LOOKUP_REVAL` set (and no
+`LOOKUP_RCU`) to ensure that entries found in the cache are forcibly
+revalidated - normally entries are only revalidated if the filesystem
+determines that they are too old to trust.
+
+The `LOOKUP_RCU` attempt may drop that flag internally and switch to
+REF-walk, but will never then try to switch back to RCU-walk.  Places
+that trip up RCU-walk are much more likely to be near the leaves and
+so it is very unlikely that there will be much, if any, benefit from
+switching back.
+
+RCU and seqlocks: fast and light
+--------------------------------
+
+RCU is, unsurprisingly, critical to RCU-walk mode.  The
+`rcu_read_lock()` is held for the entire time that RCU-walk is walking
+down a path.  The particular guarantee it provides is that the key
+data structures - dentries, inodes, super_blocks, and mounts - will
+not be freed while the lock is held.  They might be unlinked or
+invalidated in one way or another, but the memory will not be
+repurposed so values in various fields will still be meaningful.  This
+is the only guarantee that RCU provides; everything else is done using
+seqlocks.
+
+As we saw above, REF-walk holds a counted reference to the current
+dentry and the current vfsmount, and does not release those references
+before taking references to the "next" dentry or vfsmount.  It also
+sometimes takes the `d_lock` spinlock.  These references and locks are
+taken to prevent certain changes from happening.  RCU-walk must not
+take those references or locks and so cannot prevent such changes.
+Instead, it checks to see if a change has been made, and aborts or
+retries if it has.
+
+To preserve the invariant mentioned above (that RCU-walk may only make
+decisions that REF-walk could have made), it must make the checks at
+or near the same places that REF-walk holds the references.  So, when
+REF-walk increments a reference count or takes a spinlock, RCU-walk
+samples the status of a seqlock using `read_seqcount_begin()` or a
+similar function.  When REF-walk decrements the count or drops the
+lock, RCU-walk checks if the sampled status is still valid using
+`read_seqcount_retry()` or similar.
+
+However, there is a little bit more to seqlocks than that.  If
+RCU-walk accesses two different fields in a seqlock-protected
+structure, or accesses the same field twice, there is no a priori
+guarantee of any consistency between those accesses.  When consistency
+is needed - which it usually is - RCU-walk must take a copy and then
+use `read_seqcount_retry()` to validate that copy.
+
+`read_seqcount_retry()` not only checks the sequence number, but also
+imposes a memory barrier so that no memory-read instruction from
+*before* the call can be delayed until *after* the call, either by the
+CPU or by the compiler.  A simple example of this can be seen in
+`slow_dentry_cmp()` which, for filesystems which do not use simple
+byte-wise name equality, calls into the filesystem to compare a name
+against a dentry.  The length and name pointer are copied into local
+variables, then `read_seqcount_retry()` is called to confirm the two
+are consistent, and only then is `->d_compare()` called.  When
+standard filename comparison is used, `dentry_cmp()` is called
+instead.  Notably it does _not_ use `read_seqcount_retry()`, but
+instead has a large comment explaining why the consistency guarantee
+isn't necessary.  A subsequent `read_seqcount_retry()` will be
+sufficient to catch any problem that could occur at this point.
+
+With that little refresher on seqlocks out of the way we can look at
+the bigger picture of how RCU-walk uses seqlocks.
+
+### `mount_lock` and `nd->m_seq` ###
+
+We already met the `mount_lock` seqlock when REF-walk used it to
+ensure that crossing a mount point is performed safely.  RCU-walk uses
+it for that too, but for quite a bit more.
+
+Instead of taking a counted reference to each `vfsmount` as it
+descends the tree, RCU-walk samples the state of `mount_lock` at the
+start of the walk and stores this initial sequence number in the
+`struct nameidata` in the `m_seq` field.  This one lock and one
+sequence number are used to validate all accesses to all `vfsmounts`,
+and all mount point crossings.  As changes to the mount table are
+relatively rare, it is reasonable to fall back on REF-walk any time
+that any "mount" or "unmount" happens.
+
+`m_seq` is checked (using `read_seqretry()`) at the end of an RCU-walk
+sequence, whether switching to REF-walk for the rest of the path or
+when the end of the path is reached.  It is also checked when stepping
+down over a mount point (in `__follow_mount_rcu()`) or up (in
+`follow_dotdot_rcu()`).  If it is ever found to have changed, the
+whole RCU-walk sequence is aborted and the path is processed again by
+REF-walk.
+
+If RCU-walk finds that `mount_lock` hasn't changed then it can be sure
+that, had REF-walk taken counted references on each vfsmount, the
+results would have been the same.  This ensures the invariant holds,
+at least for vfsmount structures.
+
+### `dentry->d_seq` and `nd->seq`. ###
+
+In place of taking a count or lock on `d_reflock`, RCU-walk samples
+the per-dentry `d_seq` seqlock, and stores the sequence number in the
+`seq` field of the nameidata structure, so `nd->seq` should always be
+the current sequence number of `nd->dentry`.  This number needs to be
+revalidated after copying, and before using, the name, parent, or
+inode of the dentry.
+
+The handling of the name we have already looked at, and the parent is
+only accessed in `follow_dotdot_rcu()` which fairly trivially follows
+the required pattern, though it does so for three different cases.
+
+When not at a mount point, `d_parent` is followed and its `d_seq` is
+collected.  When we are at a mount point, we instead follow the
+`mnt->mnt_mountpoint` link to get a new dentry and collect its
+`d_seq`.  Then, after finally finding a `d_parent` to follow, we must
+check if we have landed on a mount point and, if so, must find that
+mount point and follow the `mnt->mnt_root` link.  This would imply a
+somewhat unusual, but certainly possible, circumstance where the
+starting point of the path lookup was in part of the filesystem that
+was mounted on, and so not visible from the root.
+
+The inode pointer, stored in `->d_inode`, is a little more
+interesting.  The inode will always need to be accessed at least
+twice, once to determine if it is NULL and once to verify access
+permissions.  Symlink handling requires a validated inode pointer too.
+Rather than revalidating on each access, a copy is made on the first
+access and it is stored in the `inode` field of `nameidata` from where
+it can be safely accessed without further validation.
+
+`lookup_fast()` is the only lookup routine that is used in RCU-mode,
+`lookup_slow()` being too slow and requiring locks.  It is in
+`lookup_fast()` that we find the important "hand over hand" tracking
+of the current dentry.
+
+The current `dentry` and current `seq` number are passed to
+`__d_lookup_rcu()` which, on success, returns a new `dentry` and a
+new `seq` number.  `lookup_fast()` then copies the inode pointer and
+revalidates the new `seq` number.  It then validates the old `dentry`
+with the old `seq` number one last time and only then continues.  This
+process of getting the `seq` number of the new dentry and then
+checking the `seq` number of the old exactly mirrors the process of
+getting a counted reference to the new dentry before dropping that for
+the old dentry which we saw in REF-walk.
+
+### No `inode->i_mutex` or even `rename_lock` ###
+
+A mutex is a fairly heavyweight lock that can only be taken when it is
+permissible to sleep.  As `rcu_read_lock()` forbids sleeping,
+`inode->i_mutex` plays no role in RCU-walk.  If some other thread does
+take `i_mutex` and modifies the directory in a way that RCU-walk needs
+to notice, the result will be either that RCU-walk fails to find the
+dentry that it is looking for, or it will find a dentry which
+`read_seqretry()` won't validate.  In either case it will drop down to
+REF-walk mode which can take whatever locks are needed.
+
+Though `rename_lock` could be used by RCU-walk as it doesn't require
+any sleeping, RCU-walk doesn't bother.  REF-walk uses `rename_lock` to
+protect against the possibility of hash chains in the dcache changing
+while they are being searched.  This can result in failing to find
+something that actually is there.  When RCU-walk fails to find
+something in the dentry cache, whether it is really there or not, it
+already drops down to REF-walk and tries again with appropriate
+locking.  This neatly handles all cases, so adding extra checks on
+rename_lock would bring no significant value.
+
+`unlazy walk()` and `complete_walk()`
+-------------------------------------
+
+That "dropping down to REF-walk" typically involves a call to
+`unlazy_walk()`, so named because "RCU-walk" is also sometimes
+referred to as "lazy walk".  `unlazy_walk()` is called when
+following the path down to the current vfsmount/dentry pair seems to
+have proceeded successfully, but the next step is problematic.  This
+can happen if the next name cannot be found in the dcache, if
+permission checking or name revalidation couldn't be achieved while
+the `rcu_read_lock()` is held (which forbids sleeping), if an
+automount point is found, or in a couple of cases involving symlinks.
+It is also called from `complete_walk()` when the lookup has reached
+the final component, or the very end of the path, depending on which
+particular flavor of lookup is used.
+
+Other reasons for dropping out of RCU-walk that do not trigger a call
+to `unlazy_walk()` are when some inconsistency is found that cannot be
+handled immediately, such as `mount_lock` or one of the `d_seq`
+seqlocks reporting a change.  In these cases the relevant function
+will return `-ECHILD` which will percolate up until it triggers a new
+attempt from the top using REF-walk.
+
+For those cases where `unlazy_walk()` is an option, it essentially
+takes a reference on each of the pointers that it holds (vfsmount,
+dentry, and possibly some symbolic links) and then verifies that the
+relevant seqlocks have not been changed.  If there have been changes,
+it, too, aborts with `-ECHILD`, otherwise the transition to REF-walk
+has been a success and the lookup process continues.
+
+Taking a reference on those pointers is not quite as simple as just
+incrementing a counter.  That works to take a second reference if you
+already have one (often indirectly through another object), but it
+isn't sufficient if you don't actually have a counted reference at
+all.  For `dentry->d_lockref`, it is safe to increment the reference
+counter to get a reference unless it has been explicitly marked as
+"dead" which involves setting the counter to `-128`.
+`lockref_get_not_dead()` achieves this.
+
+For `mnt->mnt_count` it is safe to take a reference as long as
+`mount_lock` is then used to validate the reference.  If that
+validation fails, it may *not* be safe to just drop that reference in
+the standard way of calling `mnt_put()` - an unmount may have
+progressed too far.  So the code in `legitimize_mnt()`, when it
+finds that the reference it got might not be safe, checks the
+`MNT_SYNC_UMOUNT` flag to determine if a simple `mnt_put()` is
+correct, or if it should just decrement the count and pretend none of
+this ever happened.
+
+Taking care in filesystems
+---------------------------
+
+RCU-walk depends almost entirely on cached information and often will
+not call into the filesystem at all.  However there are two places,
+besides the already-mentioned component-name comparison, where the
+file system might be included in RCU-walk, and it must know to be
+careful.
+
+If the filesystem has non-standard permission-checking requirements -
+such as a networked filesystem which may need to check with the server
+- the `i_op->permission` interface might be called during RCU-walk.
+In this case an extra "`MAY_NOT_BLOCK`" flag is passed so that it
+knows not to sleep, but to return `-ECHILD` if it cannot complete
+promptly.  `i_op->permission` is given the inode pointer, not the
+dentry, so it doesn't need to worry about further consistency checks.
+However if it accesses any other filesystem data structures, it must
+ensure they are safe to be accessed with only the `rcu_read_lock()`
+held.  This typically means they must be freed using `kfree_rcu()` or
+similar.
+
+[`READ_ONCE()`]: https://lwn.net/Articles/624126/
+
+If the filesystem may need to revalidate dcache entries, then
+`d_op->d_revalidate` may be called in RCU-walk too.  This interface
+*is* passed the dentry but does not have access to the `inode` or the
+`seq` number from the `nameidata`, so it needs to be extra careful
+when accessing fields in the dentry.  This "extra care" typically
+involves using [`READ_ONCE()`] to access fields, and verifying the
+result is not NULL before using it.  This pattern can be seen in
+`nfs_lookup_revalidate()`.
+
+A pair of patterns
+------------------
+
+In various places in the details of REF-walk and RCU-walk, and also in
+the big picture, there are a couple of related patterns that are worth
+being aware of.
+
+The first is "try quickly and check, if that fails try slowly".  We
+can see that in the high-level approach of first trying RCU-walk and
+then trying REF-walk, and in places where `unlazy_walk()` is used to
+switch to REF-walk for the rest of the path.  We also saw it earlier
+in `dget_parent()` when following a "`..`" link.  It tries a quick way
+to get a reference, then falls back to taking locks if needed.
+
+The second pattern is "try quickly and check, if that fails try
+again - repeatedly".  This is seen with the use of `rename_lock` and
+`mount_lock` in REF-walk.  RCU-walk doesn't make use of this pattern -
+if anything goes wrong it is much safer to just abort and try a more
+sedate approach.
+
+The emphasis here is "try quickly and check".  It should probably be
+"try quickly _and carefully,_ then check".  The fact that checking is
+needed is a reminder that the system is dynamic and only a limited
+number of things are safe at all.  The most likely cause of errors in
+this whole process is assuming something is safe when in reality it
+isn't.  Careful consideration of what exactly guarantees the safety of
+each access is sometimes necessary.
+
+A walk among the symlinks
+=========================
+
+There are several basic issues that we will examine to understand the
+handling of symbolic links:  the symlink stack, together with cache
+lifetimes, will help us understand the overall recursive handling of
+symlinks and lead to the special care needed for the final component.
+Then a consideration of access-time updates and summary of the various
+flags controlling lookup will finish the story.
+
+The symlink stack
+-----------------
+
+There are only two sorts of filesystem objects that can usefully
+appear in a path prior to the final component: directories and symlinks.
+Handling directories is quite straightforward: the new directory
+simply becomes the starting point at which to interpret the next
+component on the path.  Handling symbolic links requires a bit more
+work.
+
+Conceptually, symbolic links could be handled by editing the path.  If
+a component name refers to a symbolic link, then that component is
+replaced by the body of the link and, if that body starts with a '/',
+then all preceding parts of the path are discarded.  This is what the
+"`readlink -f`" command does, though it also edits out "`.`" and
+"`..`" components.
+
+Directly editing the path string is not really necessary when looking
+up a path, and discarding early components is pointless as they aren't
+looked at anyway.  Keeping track of all remaining components is
+important, but they can of course be kept separately; there is no need
+to concatenate them.  As one symlink may easily refer to another,
+which in turn can refer to a third, we may need to keep the remaining
+components of several paths, each to be processed when the preceding
+ones are completed.  These path remnants are kept on a stack of
+limited size.
+
+There are two reasons for placing limits on how many symlinks can
+occur in a single path lookup.  The most obvious is to avoid loops.
+If a symlink referred to itself either directly or through
+intermediaries, then following the symlink can never complete
+successfully - the error `ELOOP` must be returned.  Loops can be
+detected without imposing limits, but limits are the simplest solution
+and, given the second reason for restriction, quite sufficient.
+
+[outlined recently]: http://thread.gmane.org/gmane.linux.kernel/1934390/focus=1934550
+
+The second reason was [outlined recently] by Linus:
+
+>  Because it's a latency and DoS issue too. We need to react well to
+>  true loops, but also to "very deep" non-loops. It's not about memory
+>  use, it's about users triggering unreasonable CPU resources.
+
+Linux imposes a limit on the length of any pathname: `PATH_MAX`, which
+is 4096.  There are a number of reasons for this limit; not letting the
+kernel spend too much time on just one path is one of them.  With
+symbolic links you can effectively generate much longer paths so some
+sort of limit is needed for the same reason.  Linux imposes a limit of
+at most 40 symlinks in any one path lookup.  It previously imposed a
+further limit of eight on the maximum depth of recursion, but that was
+raised to 40 when a separate stack was implemented, so there is now
+just the one limit.
+
+The `nameidata` structure that we met in an earlier article contains a
+small stack that can be used to store the remaining part of up to two
+symlinks.  In many cases this will be sufficient.  If it isn't, a
+separate stack is allocated with room for 40 symlinks.  Pathname
+lookup will never exceed that stack as, once the 40th symlink is
+detected, an error is returned.
+
+It might seem that the name remnants are all that needs to be stored on
+this stack, but we need a bit more.  To see that, we need to move on to
+cache lifetimes.
+
+Storage and lifetime of cached symlinks
+---------------------------------------
+
+Like other filesystem resources, such as inodes and directory
+entries, symlinks are cached by Linux to avoid repeated costly access
+to external storage.  It is particularly important for RCU-walk to be
+able to find and temporarily hold onto these cached entries, so that
+it doesn't need to drop down into REF-walk.
+
+[object-oriented design pattern]: https://lwn.net/Articles/446317/
+
+While each filesystem is free to make its own choice, symlinks are
+typically stored in one of two places.  Short symlinks are often
+stored directly in the inode.  When a filesystem allocates a `struct
+inode` it typically allocates extra space to store private data (a
+common [object-oriented design pattern] in the kernel).  This will
+sometimes include space for a symlink.  The other common location is
+in the page cache, which normally stores the content of files.  The
+pathname in a symlink can be seen as the content of that symlink and
+can easily be stored in the page cache just like file content.
+
+When neither of these is suitable, the next most likely scenario is
+that the filesystem will allocate some temporary memory and copy or
+construct the symlink content into that memory whenever it is needed.
+
+When the symlink is stored in the inode, it has the same lifetime as
+the inode which, itself, is protected by RCU or by a counted reference
+on the dentry.  This means that the mechanisms that pathname lookup
+uses to access the dcache and icache (inode cache) safely are quite
+sufficient for accessing some cached symlinks safely.  In these cases,
+the `i_link` pointer in the inode is set to point to wherever the
+symlink is stored and it can be accessed directly whenever needed.
+
+When the symlink is stored in the page cache or elsewhere, the
+situation is not so straightforward.  A reference on a dentry or even
+on an inode does not imply any reference on cached pages of that
+inode, and even an `rcu_read_lock()` is not sufficient to ensure that
+a page will not disappear.  So for these symlinks the pathname lookup
+code needs to ask the filesystem to provide a stable reference and,
+significantly, needs to release that reference when it is finished
+with it.
+
+Taking a reference to a cache page is often possible even in RCU-walk
+mode.  It does require making changes to memory, which is best avoided,
+but that isn't necessarily a big cost and it is better than dropping
+out of RCU-walk mode completely.  Even filesystems that allocate
+space to copy the symlink into can use `GFP_ATOMIC` to often successfully
+allocate memory without the need to drop out of RCU-walk.  If a
+filesystem cannot successfully get a reference in RCU-walk mode, it
+must return `-ECHILD` and `unlazy_walk()` will be called to return to
+REF-walk mode in which the filesystem is allowed to sleep.
+
+The place for all this to happen is the `i_op->follow_link()` inode
+method.  In the present mainline code this is never actually called in
+RCU-walk mode as the rewrite is not quite complete.  It is likely that
+in a future release this method will be passed an `inode` pointer when
+called in RCU-walk mode so it both (1) knows to be careful, and (2) has the
+validated pointer.  Much like the `i_op->permission()` method we
+looked at previously, `->follow_link()` would need to be careful that
+all the data structures it references are safe to be accessed while
+holding no counted reference, only the RCU lock.  Though getting a
+reference with `->follow_link()` is not yet done in RCU-walk mode, the
+code is ready to release the reference when that does happen.
+
+This need to drop the reference to a symlink adds significant
+complexity.  It requires a reference to the inode so that the
+`i_op->put_link()` inode operation can be called.  In REF-walk, that
+reference is kept implicitly through a reference to the dentry, so
+keeping the `struct path` of the symlink is easiest.  For RCU-walk,
+the pointer to the inode is kept separately.  To allow switching from
+RCU-walk back to REF-walk in the middle of processing nested symlinks
+we also need the seq number for the dentry so we can confirm that
+switching back was safe.
+
+Finally, when providing a reference to a symlink, the filesystem also
+provides an opaque "cookie" that must be passed to `->put_link()` so that it
+knows what to free.  This might be the allocated memory area, or a
+pointer to the `struct page` in the page cache, or something else
+completely.  Only the filesystem knows what it is.
+
+In order for the reference to each symlink to be dropped when the walk completes,
+whether in RCU-walk or REF-walk, the symlink stack needs to contain,
+along with the path remnants:
+
+- the `struct path` to provide a reference to the inode in REF-walk
+- the `struct inode *` to provide a reference to the inode in RCU-walk
+- the `seq` to allow the path to be safely switched from RCU-walk to REF-walk
+- the `cookie` that tells `->put_path()` what to put.
+
+This means that each entry in the symlink stack needs to hold five
+pointers and an integer instead of just one pointer (the path
+remnant).  On a 64-bit system, this is about 40 bytes per entry;
+with 40 entries it adds up to 1600 bytes total, which is less than
+half a page.  So it might seem like a lot, but is by no means
+excessive.
+
+Note that, in a given stack frame, the path remnant (`name`) is not
+part of the symlink that the other fields refer to.  It is the remnant
+to be followed once that symlink has been fully parsed.
+
+Following the symlink
+---------------------
+
+The main loop in `link_path_walk()` iterates seamlessly over all
+components in the path and all of the non-final symlinks.  As symlinks
+are processed, the `name` pointer is adjusted to point to a new
+symlink, or is restored from the stack, so that much of the loop
+doesn't need to notice.  Getting this `name` variable on and off the
+stack is very straightforward; pushing and popping the references is
+a little more complex.
+
+When a symlink is found, `walk_component()` returns the value `1`
+(`0` is returned for any other sort of success, and a negative number
+is, as usual, an error indicator).  This causes `get_link()` to be
+called; it then gets the link from the filesystem.  Providing that
+operation is successful, the old path `name` is placed on the stack,
+and the new value is used as the `name` for a while.  When the end of
+the path is found (i.e. `*name` is `'\0'`) the old `name` is restored
+off the stack and path walking continues.
+
+Pushing and popping the reference pointers (inode, cookie, etc.) is more
+complex in part because of the desire to handle tail recursion.  When
+the last component of a symlink itself points to a symlink, we
+want to pop the symlink-just-completed off the stack before pushing
+the symlink-just-found to avoid leaving empty path remnants that would
+just get in the way.
+
+It is most convenient to push the new symlink references onto the
+stack in `walk_component()` immediately when the symlink is found;
+`walk_component()` is also the last piece of code that needs to look at the
+old symlink as it walks that last component.  So it is quite
+convenient for `walk_component()` to release the old symlink and pop
+the references just before pushing the reference information for the
+new symlink.  It is guided in this by two flags; `WALK_GET`, which
+gives it permission to follow a symlink if it finds one, and
+`WALK_PUT`, which tells it to release the current symlink after it has been
+followed.  `WALK_PUT` is tested first, leading to a call to
+`put_link()`.  `WALK_GET` is tested subsequently (by
+`should_follow_link()`) leading to a call to `pick_link()` which sets
+up the stack frame.
+
+### Symlinks with no final component ###
+
+A pair of special-case symlinks deserve a little further explanation.
+Both result in a new `struct path` (with mount and dentry) being set
+up in the `nameidata`, and result in `get_link()` returning `NULL`.
+
+The more obvious case is a symlink to "`/`".  All symlinks starting
+with "`/`" are detected in `get_link()` which resets the `nameidata`
+to point to the effective filesystem root.  If the symlink only
+contains "`/`" then there is nothing more to do, no components at all,
+so `NULL` is returned to indicate that the symlink can be released and
+the stack frame discarded.
+
+The other case involves things in `/proc` that look like symlinks but
+aren't really.
+
+>     $ ls -l /proc/self/fd/1
+>     lrwx------ 1 neilb neilb 64 Jun 13 10:19 /proc/self/fd/1 -> /dev/pts/4
+
+Every open file descriptor in any process is represented in `/proc` by
+something that looks like a symlink.  It is really a reference to the
+target file, not just the name of it.  When you `readlink` these
+objects you get a name that might refer to the same file - unless it
+has been unlinked or mounted over.  When `walk_component()` follows
+one of these, the `->follow_link()` method in "procfs" doesn't return
+a string name, but instead calls `nd_jump_link()` which updates the
+`nameidata` in place to point to that target.  `->follow_link()` then
+returns `NULL`.  Again there is no final component and `get_link()`
+reports this by leaving the `last_type` field of `nameidata` as
+`LAST_BIND`.
+
+Following the symlink in the final component
+--------------------------------------------
+
+All this leads to `link_path_walk()` walking down every component, and
+following all symbolic links it finds, until it reaches the final
+component.  This is just returned in the `last` field of `nameidata`.
+For some callers, this is all they need; they want to create that
+`last` name if it doesn't exist or give an error if it does.  Other
+callers will want to follow a symlink if one is found, and possibly
+apply special handling to the last component of that symlink, rather
+than just the last component of the original file name.  These callers
+potentially need to call `link_path_walk()` again and again on
+successive symlinks until one is found that doesn't point to another
+symlink.
+
+This case is handled by the relevant caller of `link_path_walk()`, such as
+`path_lookupat()` using a loop that calls `link_path_walk()`, and then
+handles the final component.  If the final component is a symlink
+that needs to be followed, then `trailing_symlink()` is called to set
+things up properly and the loop repeats, calling `link_path_walk()`
+again.  This could loop as many as 40 times if the last component of
+each symlink is another symlink.
+
+The various functions that examine the final component and possibly
+report that it is a symlink are `lookup_last()`, `mountpoint_last()`
+and `do_last()`, each of which use the same convention as
+`walk_component()` of returning `1` if a symlink was found that needs
+to be followed.
+
+Of these, `do_last()` is the most interesting as it is used for
+opening a file.  Part of `do_last()` runs with `i_mutex` held and this
+part is in a separate function: `lookup_open()`.
+
+Explaining `do_last()` completely is beyond the scope of this article,
+but a few highlights should help those interested in exploring the
+code.
+
+1. Rather than just finding the target file, `do_last()` needs to open
+ it.  If the file was found in the dcache, then `vfs_open()` is used for
+ this.  If not, then `lookup_open()` will either call `atomic_open()` (if
+ the filesystem provides it) to combine the final lookup with the open, or
+ will perform the separate `lookup_real()` and `vfs_create()` steps
+ directly.  In the later case the actual "open" of this newly found or
+ created file will be performed by `vfs_open()`, just as if the name
+ were found in the dcache.
+
+2. `vfs_open()` can fail with `-EOPENSTALE` if the cached information
+ wasn't quite current enough.  Rather than restarting the lookup from
+ the top with `LOOKUP_REVAL` set, `lookup_open()` is called instead,
+ giving the filesystem a chance to resolve small inconsistencies.
+ If that doesn't work, only then is the lookup restarted from the top.
+
+3. An open with O_CREAT **does** follow a symlink in the final component,
+     unlike other creation system calls (like `mkdir`).  So the sequence:
+
+     >     ln -s bar /tmp/foo
+     >     echo hello > /tmp/foo
+
+     will create a file called `/tmp/bar`.  This is not permitted if
+     `O_EXCL` is set but otherwise is handled for an O_CREAT open much
+     like for a non-creating open: `should_follow_link()` returns `1`, and
+     so does `do_last()` so that `trailing_symlink()` gets called and the
+     open process continues on the symlink that was found.
+
+Updating the access time
+------------------------
+
+We previously said of RCU-walk that it would "take no locks, increment
+no counts, leave no footprints."  We have since seen that some
+"footprints" can be needed when handling symlinks as a counted
+reference (or even a memory allocation) may be needed.  But these
+footprints are best kept to a minimum.
+
+One other place where walking down a symlink can involve leaving
+footprints in a way that doesn't affect directories is in updating access times.
+In Unix (and Linux) every filesystem object has a "last accessed
+time", or "`atime`".  Passing through a directory to access a file
+within is not considered to be an access for the purposes of
+`atime`; only listing the contents of a directory can update its `atime`.
+Symlinks are different it seems.  Both reading a symlink (with `readlink()`)
+and looking up a symlink on the way to some other destination can
+update the atime on that symlink.
+
+[clearest statement]: http://pubs.opengroup.org/onlinepubs/9699919799/basedefs/V1_chap04.html#tag_04_08
+
+It is not clear why this is the case; POSIX has little to say on the
+subject.  The [clearest statement] is that, if a particular implementation
+updates a timestamp in a place not specified by POSIX, this must be
+documented "except that any changes caused by pathname resolution need
+not be documented".  This seems to imply that POSIX doesn't really
+care about access-time updates during pathname lookup.
+
+[Linux 1.3.87]: https://git.kernel.org/cgit/linux/kernel/git/history/history.git/diff/fs/ext2/symlink.c?id=f806c6db77b8eaa6e00dcfb6b567706feae8dbb8
+
+An examination of history shows that prior to [Linux 1.3.87], the ext2
+filesystem, at least, didn't update atime when following a link.
+Unfortunately we have no record of why that behavior was changed.
+
+In any case, access time must now be updated and that operation can be
+quite complex.  Trying to stay in RCU-walk while doing it is best
+avoided.  Fortunately it is often permitted to skip the `atime`
+update.  Because `atime` updates cause performance problems in various
+areas, Linux supports the `relatime` mount option, which generally
+limits the updates of `atime` to once per day on files that aren't
+being changed (and symlinks never change once created).  Even without
+`relatime`, many filesystems record `atime` with a one-second
+granularity, so only one update per second is required.
+
+It is easy to test if an `atime` update is needed while in RCU-walk
+mode and, if it isn't, the update can be skipped and RCU-walk mode
+continues.  Only when an `atime` update is actually required does the
+path walk drop down to REF-walk.  All of this is handled in the
+`get_link()` function.
+
+A few flags
+-----------
+
+A suitable way to wrap up this tour of pathname walking is to list
+the various flags that can be stored in the `nameidata` to guide the
+lookup process.  Many of these are only meaningful on the final
+component, others reflect the current state of the pathname lookup.
+And then there is `LOOKUP_EMPTY`, which doesn't fit conceptually with
+the others.  If this is not set, an empty pathname causes an error
+very early on.  If it is set, empty pathnames are not considered to be
+an error.
+
+### Global state flags ###
+
+We have already met two global state flags: `LOOKUP_RCU` and
+`LOOKUP_REVAL`.  These select between one of three overall approaches
+to lookup: RCU-walk, REF-walk, and REF-walk with forced revalidation.
+
+`LOOKUP_PARENT` indicates that the final component hasn't been reached
+yet.  This is primarily used to tell the audit subsystem the full
+context of a particular access being audited.
+
+`LOOKUP_ROOT` indicates that the `root` field in the `nameidata` was
+provided by the caller, so it shouldn't be released when it is no
+longer needed.
+
+`LOOKUP_JUMPED` means that the current dentry was chosen not because
+it had the right name but for some other reason.  This happens when
+following "`..`", following a symlink to `/`, crossing a mount point
+or accessing a "`/proc/$PID/fd/$FD`" symlink.  In this case the
+filesystem has not been asked to revalidate the name (with
+`d_revalidate()`).  In such cases the inode may still need to be
+revalidated, so `d_op->d_weak_revalidate()` is called if
+`LOOKUP_JUMPED` is set when the look completes - which may be at the
+final component or, when creating, unlinking, or renaming, at the penultimate component.
+
+### Final-component flags ###
+
+Some of these flags are only set when the final component is being
+considered.  Others are only checked for when considering that final
+component.
+
+`LOOKUP_AUTOMOUNT` ensures that, if the final component is an automount
+point, then the mount is triggered.  Some operations would trigger it
+anyway, but operations like `stat()` deliberately don't.  `statfs()`
+needs to trigger the mount but otherwise behaves a lot like `stat()`, so
+it sets `LOOKUP_AUTOMOUNT`, as does "`quotactl()`" and the handling of
+"`mount --bind`".
+
+`LOOKUP_FOLLOW` has a similar function to `LOOKUP_AUTOMOUNT` but for
+symlinks.  Some system calls set or clear it implicitly, while
+others have API flags such as `AT_SYMLINK_FOLLOW` and
+`UMOUNT_NOFOLLOW` to control it.  Its effect is similar to
+`WALK_GET` that we already met, but it is used in a different way.
+
+`LOOKUP_DIRECTORY` insists that the final component is a directory.
+Various callers set this and it is also set when the final component
+is found to be followed by a slash.
+
+Finally `LOOKUP_OPEN`, `LOOKUP_CREATE`, `LOOKUP_EXCL`, and
+`LOOKUP_RENAME_TARGET` are not used directly by the VFS but are made
+available to the filesystem and particularly the `->d_revalidate()`
+method.  A filesystem can choose not to bother revalidating too hard
+if it knows that it will be asked to open or create the file soon.
+These flags were previously useful for `->lookup()` too but with the
+introduction of `->atomic_open()` they are less relevant there.
+
+End of the road
+---------------
+
+Despite its complexity, all this pathname lookup code appears to be
+in good shape - various parts are certainly easier to understand now
+than even a couple of releases ago.  But that doesn't mean it is
+"finished".   As already mentioned, RCU-walk currently only follows
+symlinks that are stored in the inode so, while it handles many ext4
+symlinks, it doesn't help with NFS, XFS, or Btrfs.  That support
+is not likely to be long delayed.
diff --git a/Documentation/filesystems/path-lookup.txt b/Documentation/filesystems/path-lookup.txt
new file mode 100644
index 000000000..9b8930f58
--- /dev/null
+++ b/Documentation/filesystems/path-lookup.txt
@@ -0,0 +1,382 @@
+Path walking and name lookup locking
+====================================
+
+Path resolution is the finding a dentry corresponding to a path name string, by
+performing a path walk. Typically, for every open(), stat() etc., the path name
+will be resolved. Paths are resolved by walking the namespace tree, starting
+with the first component of the pathname (eg. root or cwd) with a known dentry,
+then finding the child of that dentry, which is named the next component in the
+path string. Then repeating the lookup from the child dentry and finding its
+child with the next element, and so on.
+
+Since it is a frequent operation for workloads like multiuser environments and
+web servers, it is important to optimize this code.
+
+Path walking synchronisation history:
+Prior to 2.5.10, dcache_lock was acquired in d_lookup (dcache hash lookup) and
+thus in every component during path look-up. Since 2.5.10 onwards, fast-walk
+algorithm changed this by holding the dcache_lock at the beginning and walking
+as many cached path component dentries as possible. This significantly
+decreases the number of acquisition of dcache_lock. However it also increases
+the lock hold time significantly and affects performance in large SMP machines.
+Since 2.5.62 kernel, dcache has been using a new locking model that uses RCU to
+make dcache look-up lock-free.
+
+All the above algorithms required taking a lock and reference count on the
+dentry that was looked up, so that may be used as the basis for walking the
+next path element. This is inefficient and unscalable. It is inefficient
+because of the locks and atomic operations required for every dentry element
+slows things down. It is not scalable because many parallel applications that
+are path-walk intensive tend to do path lookups starting from a common dentry
+(usually, the root "/" or current working directory). So contention on these
+common path elements causes lock and cacheline queueing.
+
+Since 2.6.38, RCU is used to make a significant part of the entire path walk
+(including dcache look-up) completely "store-free" (so, no locks, atomics, or
+even stores into cachelines of common dentries). This is known as "rcu-walk"
+path walking.
+
+Path walking overview
+=====================
+
+A name string specifies a start (root directory, cwd, fd-relative) and a
+sequence of elements (directory entry names), which together refer to a path in
+the namespace. A path is represented as a (dentry, vfsmount) tuple. The name
+elements are sub-strings, separated by '/'.
+
+Name lookups will want to find a particular path that a name string refers to
+(usually the final element, or parent of final element). This is done by taking
+the path given by the name's starting point (which we know in advance -- eg.
+current->fs->cwd or current->fs->root) as the first parent of the lookup. Then
+iteratively for each subsequent name element, look up the child of the current
+parent with the given name and if it is not the desired entry, make it the
+parent for the next lookup.
+
+A parent, of course, must be a directory, and we must have appropriate
+permissions on the parent inode to be able to walk into it.
+
+Turning the child into a parent for the next lookup requires more checks and
+procedures. Symlinks essentially substitute the symlink name for the target
+name in the name string, and require some recursive path walking.  Mount points
+must be followed into (thus changing the vfsmount that subsequent path elements
+refer to), switching from the mount point path to the root of the particular
+mounted vfsmount. These behaviours are variously modified depending on the
+exact path walking flags.
+
+Path walking then must, broadly, do several particular things:
+- find the start point of the walk;
+- perform permissions and validity checks on inodes;
+- perform dcache hash name lookups on (parent, name element) tuples;
+- traverse mount points;
+- traverse symlinks;
+- lookup and create missing parts of the path on demand.
+
+Safe store-free look-up of dcache hash table
+============================================
+
+Dcache name lookup
+------------------
+In order to lookup a dcache (parent, name) tuple, we take a hash on the tuple
+and use that to select a bucket in the dcache-hash table. The list of entries
+in that bucket is then walked, and we do a full comparison of each entry
+against our (parent, name) tuple.
+
+The hash lists are RCU protected, so list walking is not serialised with
+concurrent updates (insertion, deletion from the hash). This is a standard RCU
+list application with the exception of renames, which will be covered below.
+
+Parent and name members of a dentry, as well as its membership in the dcache
+hash, and its inode are protected by the per-dentry d_lock spinlock. A
+reference is taken on the dentry (while the fields are verified under d_lock),
+and this stabilises its d_inode pointer and actual inode. This gives a stable
+point to perform the next step of our path walk against.
+
+These members are also protected by d_seq seqlock, although this offers
+read-only protection and no durability of results, so care must be taken when
+using d_seq for synchronisation (see seqcount based lookups, below).
+
+Renames
+-------
+Back to the rename case. In usual RCU protected lists, the only operations that
+will happen to an object is insertion, and then eventually removal from the
+list. The object will not be reused until an RCU grace period is complete.
+This ensures the RCU list traversal primitives can run over the object without
+problems (see RCU documentation for how this works).
+
+However when a dentry is renamed, its hash value can change, requiring it to be
+moved to a new hash list. Allocating and inserting a new alias would be
+expensive and also problematic for directory dentries. Latency would be far to
+high to wait for a grace period after removing the dentry and before inserting
+it in the new hash bucket. So what is done is to insert the dentry into the
+new list immediately.
+
+However, when the dentry's list pointers are updated to point to objects in the
+new list before waiting for a grace period, this can result in a concurrent RCU
+lookup of the old list veering off into the new (incorrect) list and missing
+the remaining dentries on the list.
+
+There is no fundamental problem with walking down the wrong list, because the
+dentry comparisons will never match. However it is fatal to miss a matching
+dentry. So a seqlock is used to detect when a rename has occurred, and so the
+lookup can be retried.
+
+         1      2      3
+        +---+  +---+  +---+
+hlist-->| N-+->| N-+->| N-+->
+head <--+-P |<-+-P |<-+-P |
+        +---+  +---+  +---+
+
+Rename of dentry 2 may require it deleted from the above list, and inserted
+into a new list. Deleting 2 gives the following list.
+
+         1             3
+        +---+         +---+     (don't worry, the longer pointers do not
+hlist-->| N-+-------->| N-+->    impose a measurable performance overhead
+head <--+-P |<--------+-P |      on modern CPUs)
+        +---+         +---+
+          ^      2      ^
+          |    +---+    |
+          |    | N-+----+
+          +----+-P |
+               +---+
+
+This is a standard RCU-list deletion, which leaves the deleted object's
+pointers intact, so a concurrent list walker that is currently looking at
+object 2 will correctly continue to object 3 when it is time to traverse the
+next object.
+
+However, when inserting object 2 onto a new list, we end up with this:
+
+         1             3
+        +---+         +---+
+hlist-->| N-+-------->| N-+->
+head <--+-P |<--------+-P |
+        +---+         +---+
+                 2
+               +---+
+               | N-+---->
+          <----+-P |
+               +---+
+
+Because we didn't wait for a grace period, there may be a concurrent lookup
+still at 2. Now when it follows 2's 'next' pointer, it will walk off into
+another list without ever having checked object 3.
+
+A related, but distinctly different, issue is that of rename atomicity versus
+lookup operations. If a file is renamed from 'A' to 'B', a lookup must only
+find either 'A' or 'B'. So if a lookup of 'A' returns NULL, a subsequent lookup
+of 'B' must succeed (note the reverse is not true).
+
+Between deleting the dentry from the old hash list, and inserting it on the new
+hash list, a lookup may find neither 'A' nor 'B' matching the dentry. The same
+rename seqlock is also used to cover this race in much the same way, by
+retrying a negative lookup result if a rename was in progress.
+
+Seqcount based lookups
+----------------------
+In refcount based dcache lookups, d_lock is used to serialise access to
+the dentry, stabilising it while comparing its name and parent and then
+taking a reference count (the reference count then gives a stable place to
+start the next part of the path walk from).
+
+As explained above, we would like to do path walking without taking locks or
+reference counts on intermediate dentries along the path. To do this, a per
+dentry seqlock (d_seq) is used to take a "coherent snapshot" of what the dentry
+looks like (its name, parent, and inode). That snapshot is then used to start
+the next part of the path walk. When loading the coherent snapshot under d_seq,
+care must be taken to load the members up-front, and use those pointers rather
+than reloading from the dentry later on (otherwise we'd have interesting things
+like d_inode going NULL underneath us, if the name was unlinked).
+
+Also important is to avoid performing any destructive operations (pretty much:
+no non-atomic stores to shared data), and to recheck the seqcount when we are
+"done" with the operation. Retry or abort if the seqcount does not match.
+Avoiding destructive or changing operations means we can easily unwind from
+failure.
+
+What this means is that a caller, provided they are holding RCU lock to
+protect the dentry object from disappearing, can perform a seqcount based
+lookup which does not increment the refcount on the dentry or write to
+it in any way. This returned dentry can be used for subsequent operations,
+provided that d_seq is rechecked after that operation is complete.
+
+Inodes are also rcu freed, so the seqcount lookup dentry's inode may also be
+queried for permissions.
+
+With this two parts of the puzzle, we can do path lookups without taking
+locks or refcounts on dentry elements.
+
+RCU-walk path walking design
+============================
+
+Path walking code now has two distinct modes, ref-walk and rcu-walk. ref-walk
+is the traditional[*] way of performing dcache lookups using d_lock to
+serialise concurrent modifications to the dentry and take a reference count on
+it. ref-walk is simple and obvious, and may sleep, take locks, etc while path
+walking is operating on each dentry. rcu-walk uses seqcount based dentry
+lookups, and can perform lookup of intermediate elements without any stores to
+shared data in the dentry or inode. rcu-walk can not be applied to all cases,
+eg. if the filesystem must sleep or perform non trivial operations, rcu-walk
+must be switched to ref-walk mode.
+
+[*] RCU is still used for the dentry hash lookup in ref-walk, but not the full
+    path walk.
+
+Where ref-walk uses a stable, refcounted ``parent'' to walk the remaining
+path string, rcu-walk uses a d_seq protected snapshot. When looking up a
+child of this parent snapshot, we open d_seq critical section on the child
+before closing d_seq critical section on the parent. This gives an interlocking
+ladder of snapshots to walk down.
+
+
+     proc 101
+      /----------------\
+     / comm:    "vi"    \
+    /  fs.root: dentry0  \
+    \  fs.cwd:  dentry2  /
+     \                  /
+      \----------------/
+
+So when vi wants to open("/home/npiggin/test.c", O_RDWR), then it will
+start from current->fs->root, which is a pinned dentry. Alternatively,
+"./test.c" would start from cwd; both names refer to the same path in
+the context of proc101.
+
+     dentry 0
+    +---------------------+   rcu-walk begins here, we note d_seq, check the
+    | name:    "/"        |   inode's permission, and then look up the next
+    | inode:   10         |   path element which is "home"...
+    | children:"home", ...|
+    +---------------------+
+              |
+     dentry 1 V
+    +---------------------+   ... which brings us here. We find dentry1 via
+    | name:    "home"     |   hash lookup, then note d_seq and compare name
+    | inode:   678        |   string and parent pointer. When we have a match,
+    | children:"npiggin"  |   we now recheck the d_seq of dentry0. Then we
+    +---------------------+   check inode and look up the next element.
+              |
+     dentry2  V
+    +---------------------+   Note: if dentry0 is now modified, lookup is
+    | name:    "npiggin"  |   not necessarily invalid, so we need only keep a
+    | inode:   543        |   parent for d_seq verification, and grandparents
+    | children:"a.c", ... |   can be forgotten.
+    +---------------------+
+              |
+     dentry3  V
+    +---------------------+   At this point we have our destination dentry.
+    | name:    "a.c"      |   We now take its d_lock, verify d_seq of this
+    | inode:   14221      |   dentry. If that checks out, we can increment
+    | children:NULL       |   its refcount because we're holding d_lock.
+    +---------------------+
+
+Taking a refcount on a dentry from rcu-walk mode, by taking its d_lock,
+re-checking its d_seq, and then incrementing its refcount is called
+"dropping rcu" or dropping from rcu-walk into ref-walk mode.
+
+It is, in some sense, a bit of a house of cards. If the seqcount check of the
+parent snapshot fails, the house comes down, because we had closed the d_seq
+section on the grandparent, so we have nothing left to stand on. In that case,
+the path walk must be fully restarted (which we do in ref-walk mode, to avoid
+live locks). It is costly to have a full restart, but fortunately they are
+quite rare.
+
+When we reach a point where sleeping is required, or a filesystem callout
+requires ref-walk, then instead of restarting the walk, we attempt to drop rcu
+at the last known good dentry we have. Avoiding a full restart in ref-walk in
+these cases is fundamental for performance and scalability because blocking
+operations such as creates and unlinks are not uncommon.
+
+The detailed design for rcu-walk is like this:
+* LOOKUP_RCU is set in nd->flags, which distinguishes rcu-walk from ref-walk.
+* Take the RCU lock for the entire path walk, starting with the acquiring
+  of the starting path (eg. root/cwd/fd-path). So now dentry refcounts are
+  not required for dentry persistence.
+* synchronize_rcu is called when unregistering a filesystem, so we can
+  access d_ops and i_ops during rcu-walk.
+* Similarly take the vfsmount lock for the entire path walk. So now mnt
+  refcounts are not required for persistence. Also we are free to perform mount
+  lookups, and to assume dentry mount points and mount roots are stable up and
+  down the path.
+* Have a per-dentry seqlock to protect the dentry name, parent, and inode,
+  so we can load this tuple atomically, and also check whether any of its
+  members have changed.
+* Dentry lookups (based on parent, candidate string tuple) recheck the parent
+  sequence after the child is found in case anything changed in the parent
+  during the path walk.
+* inode is also RCU protected so we can load d_inode and use the inode for
+  limited things.
+* i_mode, i_uid, i_gid can be tested for exec permissions during path walk.
+* i_op can be loaded.
+* When the destination dentry is reached, drop rcu there (ie. take d_lock,
+  verify d_seq, increment refcount).
+* If seqlock verification fails anywhere along the path, do a full restart
+  of the path lookup in ref-walk mode. -ECHILD tends to be used (for want of
+  a better errno) to signal an rcu-walk failure.
+
+The cases where rcu-walk cannot continue are:
+* NULL dentry (ie. any uncached path element)
+* Following links
+
+It may be possible eventually to make following links rcu-walk aware.
+
+Uncached path elements will always require dropping to ref-walk mode, at the
+very least because i_mutex needs to be grabbed, and objects allocated.
+
+Final note:
+"store-free" path walking is not strictly store free. We take vfsmount lock
+and refcounts (both of which can be made per-cpu), and we also store to the
+stack (which is essentially CPU-local), and we also have to take locks and
+refcount on final dentry.
+
+The point is that shared data, where practically possible, is not locked
+or stored into. The result is massive improvements in performance and
+scalability of path resolution.
+
+
+Interesting statistics
+======================
+
+The following table gives rcu lookup statistics for a few simple workloads
+(2s12c24t Westmere, debian non-graphical system). Ungraceful are attempts to
+drop rcu that fail due to d_seq failure and requiring the entire path lookup
+again. Other cases are successful rcu-drops that are required before the final
+element, nodentry for missing dentry, revalidate for filesystem revalidate
+routine requiring rcu drop, permission for permission check requiring drop,
+and link for symlink traversal requiring drop.
+
+     rcu-lookups     restart  nodentry          link  revalidate  permission
+bootup     47121           0      4624          1010       10283        7852
+dbench  25386793           0   6778659(26.7%)     55         549        1156
+kbuild   2696672          10     64442(2.3%)  108764(4.0%)     1        1590
+git diff   39605           0        28             2           0         106
+vfstest 24185492        4945    708725(2.9%) 1076136(4.4%)     0        2651
+
+What this shows is that failed rcu-walk lookups, ie. ones that are restarted
+entirely with ref-walk, are quite rare. Even the "vfstest" case which
+specifically has concurrent renames/mkdir/rmdir/ creat/unlink/etc to exercise
+such races is not showing a huge amount of restarts.
+
+Dropping from rcu-walk to ref-walk mean that we have encountered a dentry where
+the reference count needs to be taken for some reason. This is either because
+we have reached the target of the path walk, or because we have encountered a
+condition that can't be resolved in rcu-walk mode.  Ideally, we drop rcu-walk
+only when we have reached the target dentry, so the other statistics show where
+this does not happen.
+
+Note that a graceful drop from rcu-walk mode due to something such as the
+dentry not existing (which can be common) is not necessarily a failure of
+rcu-walk scheme, because some elements of the path may have been walked in
+rcu-walk mode. The further we get from common path elements (such as cwd or
+root), the less contended the dentry is likely to be. The closer we are to
+common path elements, the more likely they will exist in dentry cache.
+
+
+Papers and other documentation on dcache locking
+================================================
+
+1. Scaling dcache with RCU (http://linuxjournal.com/article.php?sid=7124).
+
+2. http://lse.sourceforge.net/locking/dcache/dcache.html
+
+3. path-lookup.md in this directory.
diff --git a/Documentation/filesystems/pohmelfs/design_notes.txt b/Documentation/filesystems/pohmelfs/design_notes.txt
new file mode 100644
index 000000000..106d17fbb
--- /dev/null
+++ b/Documentation/filesystems/pohmelfs/design_notes.txt
@@ -0,0 +1,72 @@
+POHMELFS: Parallel Optimized Host Message Exchange Layered File System.
+
+		Evgeniy Polyakov <zbr@ioremap.net>
+
+Homepage: http://www.ioremap.net/projects/pohmelfs
+
+POHMELFS first began as a network filesystem with coherent local data and
+metadata caches but is now evolving into a parallel distributed filesystem.
+
+Main features of this FS include:
+ * Locally coherent cache for data and metadata with (potentially) byte-range locks.
+	Since all Linux filesystems lock the whole inode during writing, algorithm
+	is very simple and does not use byte-ranges, although they are sent in
+	locking messages.
+ * Completely async processing of all events except creation of hard and symbolic
+	links, and rename events.
+	Object creation and data reading and writing are processed asynchronously.
+ * Flexible object architecture optimized for network processing.
+	Ability to create long paths to objects and remove arbitrarily huge
+	directories with a single network command.
+	(like removing the whole kernel tree via a single network command).
+ * Very high performance.
+ * Fast and scalable multithreaded userspace server. Being in userspace it works
+	with any underlying filesystem and still is much faster than async in-kernel NFS one.
+ * Client is able to switch between different servers (if one goes down, client
+	automatically reconnects to second and so on).
+ * Transactions support. Full failover for all operations.
+	Resending transactions to different servers on timeout or error.
+ * Read request (data read, directory listing, lookup requests) balancing between multiple servers.
+ * Write requests are replicated to multiple servers and completed only when all of them are acked.
+ * Ability to add and/or remove servers from the working set at run-time.
+ * Strong authentication and possible data encryption in network channel.
+ * Extended attributes support.
+
+POHMELFS is based on transactions, which are potentially long-standing objects that live
+in the client's memory. Each transaction contains all the information needed to process a given
+command (or set of commands, which is frequently used during data writing: single transactions
+can contain creation and data writing commands). Transactions are committed by all the servers
+to which they are sent and, in case of failures, are eventually resent or dropped with an error.
+For example, reading will return an error if no servers are available.
+
+POHMELFS uses a asynchronous approach to data processing. Courtesy of transactions, it is
+possible to detach replies from requests and, if the command requires data to be received, the
+caller sleeps waiting for it. Thus, it is possible to issue multiple read commands to different
+servers and async threads will pick up replies in parallel, find appropriate transactions in the
+system and put the data where it belongs (like the page or inode cache).
+
+The main feature of POHMELFS is writeback data and the metadata cache.
+Only a few non-performance critical operations use the write-through cache and
+are synchronous: hard and symbolic link creation, and object rename. Creation,
+removal of objects and data writing are asynchronous and are sent to
+the server during system writeback. Only one writer at a time is allowed for any
+given inode, which is guarded by an appropriate locking protocol.
+Because of this feature, POHMELFS is extremely fast at metadata intensive
+workloads and can fully utilize the bandwidth to the servers when doing bulk
+data transfers.
+
+POHMELFS clients operate with a working set of servers and are capable of balancing read-only
+operations (like lookups or directory listings) between them according to IO priorities.
+Administrators can add or remove servers from the set at run-time via special commands (described
+in Documentation/filesystems/pohmelfs/info.txt file). Writes are replicated to all servers, which
+are connected with write permission turned on. IO priority and permissions can be changed in
+run-time.
+
+POHMELFS is capable of full data channel encryption and/or strong crypto hashing.
+One can select any kernel supported cipher, encryption mode, hash type and operation mode
+(hmac or digest). It is also possible to use both or neither (default). Crypto configuration
+is checked during mount time and, if the server does not support it, appropriate capabilities
+will be disabled or mount will fail (if 'crypto_fail_unsupported' mount option is specified).
+Crypto performance heavily depends on the number of crypto threads, which asynchronously perform
+crypto operations and send the resulting data to server or submit it up the stack. This number
+can be controlled via a mount option.
diff --git a/Documentation/filesystems/pohmelfs/info.txt b/Documentation/filesystems/pohmelfs/info.txt
new file mode 100644
index 000000000..db2e41393
--- /dev/null
+++ b/Documentation/filesystems/pohmelfs/info.txt
@@ -0,0 +1,99 @@
+POHMELFS usage information.
+
+Mount options.
+All but index, number of crypto threads and maximum IO size can changed via remount.
+
+idx=%u
+ Each mountpoint is associated with a special index via this option.
+ Administrator can add or remove servers from the given index, so all mounts,
+ which were attached to it, are updated.
+ Default it is 0.
+
+trans_scan_timeout=%u
+ This timeout, expressed in milliseconds, specifies time to scan transaction
+ trees looking for stale requests, which have to be resent, or if number of
+ retries exceed specified limit, dropped with error.
+ Default is 5 seconds.
+
+drop_scan_timeout=%u
+ Internal timeout, expressed in milliseconds, which specifies how frequently
+ inodes marked to be dropped are freed. It also specifies how frequently
+ the system checks that servers have to be added or removed from current working set.
+ Default is 1 second.
+
+wait_on_page_timeout=%u
+ Number of milliseconds to wait for reply from remote server for data reading command.
+ If this timeout is exceeded, reading returns an error.
+ Default is 5 seconds.
+
+trans_retries=%u
+ This is the number of times that a transaction will be resent to a server that did
+ not answer for the last @trans_scan_timeout milliseconds.
+ When the number of resends exceeds this limit, the transaction is completed with error.
+ Default is 5 resends.
+
+crypto_thread_num=%u
+ Number of crypto processing threads. Threads are used both for RX and TX traffic.
+ Default is 2, or no threads if crypto operations are not supported.
+
+trans_max_pages=%u
+ Maximum number of pages in a single transaction. This parameter also controls
+ the number of pages,  allocated for crypto processing (each crypto thread has
+ pool of pages, the number of which is equal to 'trans_max_pages'.
+ Default is 100 pages.
+
+crypto_fail_unsupported
+ If specified, mount will fail if the server does not support requested crypto operations.
+ By default mount will disable non-matching crypto operations.
+
+mcache_timeout=%u
+ Maximum number of milliseconds to wait for the mcache objects to be processed.
+ Mcache includes locks (given lock should be granted by server), attributes (they should be
+ fully received in the given timeframe).
+ Default is 5 seconds.
+
+Usage examples.
+
+Add server server1.net:1025 into the working set with index $idx
+with appropriate hash algorithm and key file and cipher algorithm, mode and key file:
+$cfg A add -a server1.net -p 1025 -i $idx -K $hash_key -k $cipher_key
+
+Mount filesystem with given index $idx to /mnt mountpoint.
+Client will connect to all servers specified in the working set via previous command:
+mount -t pohmel -o idx=$idx q /mnt
+
+Change permissions to read-only (-I 1 option, '-I 2' - write-only, 3 - rw):
+$cfg A modify -a server1.net -p 1025 -i $idx -I 1
+
+Change IO priority to 123 (node with the highest priority gets read requests).
+$cfg A modify -a server1.net -p 1025 -i $idx -P 123
+
+One can check currect status of all connections in the mountstats file:
+# cat /proc/$PID/mountstats
+...
+device none mounted on /mnt with fstype pohmel
+idx addr(:port) socket_type protocol active priority permissions
+0 server1.net:1026 1 6 1 250 1
+0 server2.net:1025 1 6 1 123 3
+
+Server installation.
+
+Creating a server, which listens at port 1025 and 0.0.0.0 address.
+Working root directory (note, that server chroots there, so you have to have appropriate permissions)
+is set to /mnt, server will negotiate hash/cipher with client, in case client requested it, there
+are appropriate key files.
+Number of working threads is set to 10.
+
+# ./fserver -a 0.0.0.0 -p 1025 -r /mnt -w 10 -K hash_key -k cipher_key
+
+ -A 6			 - listen on ipv6 address. Default: Disabled.
+ -r root                 - path to root directory. Default: /tmp.
+ -a addr                 - listen address. Default: 0.0.0.0.
+ -p port                 - listen port. Default: 1025.
+ -w workers              - number of workers per connected client. Default: 1.
+ -K file		 - hash key size. Default: none.
+ -k file		 - cipher key size. Default: none.
+ -h                      - this help.
+
+Number of worker threads specifies how many workers will be created for each client.
+Bulk single-client transafers usually are better handled with smaller number (like 1-3).
diff --git a/Documentation/filesystems/pohmelfs/network_protocol.txt b/Documentation/filesystems/pohmelfs/network_protocol.txt
new file mode 100644
index 000000000..c680b4b53
--- /dev/null
+++ b/Documentation/filesystems/pohmelfs/network_protocol.txt
@@ -0,0 +1,227 @@
+POHMELFS network protocol.
+
+Basic structure used in network communication is following command:
+
+struct netfs_cmd
+{
+	__u16			cmd;	/* Command number */
+	__u16			csize;	/* Attached crypto information size */
+	__u16			cpad;	/* Attached padding size */
+	__u16			ext;	/* External flags */
+	__u32			size;	/* Size of the attached data */
+	__u32			trans;	/* Transaction id */
+	__u64			id;	/* Object ID to operate on. Used for feedback.*/
+	__u64			start;	/* Start of the object. */
+	__u64			iv;	/* IV sequence */
+	__u8			data[0];
+};
+
+Commands can be embedded into transaction command (which in turn has own command),
+so one can extend protocol as needed without breaking backward compatibility as long
+as old commands are supported. All string lengths include tail 0 byte.
+
+All commands are transferred over the network in big-endian. CPU endianness is used at the end peers.
+
+@cmd - command number, which specifies command to be processed. Following
+	commands are used currently:
+
+	NETFS_READDIR	= 1,	/* Read directory for given inode number */
+	NETFS_READ_PAGE,	/* Read data page from the server */
+	NETFS_WRITE_PAGE,	/* Write data page to the server */
+	NETFS_CREATE,		/* Create directory entry */
+	NETFS_REMOVE,		/* Remove directory entry */
+	NETFS_LOOKUP,		/* Lookup single object */
+	NETFS_LINK,		/* Create a link */
+	NETFS_TRANS,		/* Transaction */
+	NETFS_OPEN,		/* Open intent */
+	NETFS_INODE_INFO,	/* Metadata cache coherency synchronization message */
+	NETFS_PAGE_CACHE,	/* Page cache invalidation message */
+	NETFS_READ_PAGES,	/* Read multiple contiguous pages in one go */
+	NETFS_RENAME,		/* Rename object */
+	NETFS_CAPABILITIES,	/* Capabilities of the client, for example supported crypto */
+	NETFS_LOCK,		/* Distributed lock message */
+	NETFS_XATTR_SET,	/* Set extended attribute */
+	NETFS_XATTR_GET,	/* Get extended attribute */
+
+@ext - external flags. Used by different commands to specify some extra arguments
+	like partial size of the embedded objects or creation flags.
+
+@size - size of the attached data. For NETFS_READ_PAGE and NETFS_READ_PAGES no data is attached,
+	but size of the requested data is incorporated here. It does not include size of the command
+	header (struct netfs_cmd) itself.
+
+@id - id of the object this command operates on. Each command can use it for own purpose.
+
+@start - start of the object this command operates on. Each command can use it for own purpose.
+
+@csize, @cpad - size and padding size of the (attached if needed) crypto information.
+
+Command specifications.
+
+@NETFS_READDIR
+This command is used to sync content of the remote dir to the client.
+
+@ext - length of the path to object.
+@size - the same.
+@id - local inode number of the directory to read.
+@start - zero.
+
+
+@NETFS_READ_PAGE
+This command is used to read data from remote server.
+Data size does not exceed local page cache size.
+
+@id - inode number.
+@start - first byte offset.
+@size - number of bytes to read plus length of the path to object.
+@ext - object path length.
+
+
+@NETFS_CREATE
+Used to create object.
+It does not require that all directories on top of the object were
+already created, it will create them automatically. Each object has
+associated @netfs_path_entry data structure, which contains creation
+mode (permissions and type) and length of the name as long as name itself.
+
+@start - 0
+@size - size of the all data structures needed to create a path
+@id - local inode number
+@ext - 0
+
+
+@NETFS_REMOVE
+Used to remove object.
+
+@ext - length of the path to object.
+@size - the same.
+@id - local inode number.
+@start - zero.
+
+
+@NETFS_LOOKUP
+Lookup information about object on server.
+
+@ext - length of the path to object.
+@size - the same.
+@id - local inode number of the directory to look object in.
+@start - local inode number of the object to look at.
+
+
+@NETFS_LINK
+Create hard of symlink.
+Command is sent as "object_path|target_path".
+
+@size - size of the above string.
+@id - parent local inode number.
+@start - 1 for symlink, 0 for hardlink.
+@ext - size of the "object_path" above.
+
+
+@NETFS_TRANS
+Transaction header.
+
+@size - incorporates all embedded command sizes including theirs header sizes.
+@start - transaction generation number - unique id used to find transaction.
+@ext - transaction flags. Unused at the moment.
+@id - 0.
+
+
+@NETFS_OPEN
+Open intent for given transaction.
+
+@id - local inode number.
+@start - 0.
+@size - path length to the object.
+@ext - open flags (O_RDWR and so on).
+
+
+@NETFS_INODE_INFO
+Metadata update command.
+It is sent to servers when attributes of the object are changed and received
+when data or metadata were updated. It operates with the following structure:
+
+struct netfs_inode_info
+{
+	unsigned int		mode;
+	unsigned int		nlink;
+	unsigned int		uid;
+	unsigned int		gid;
+	unsigned int		blocksize;
+	unsigned int		padding;
+	__u64			ino;
+	__u64			blocks;
+	__u64			rdev;
+	__u64			size;
+	__u64			version;
+};
+
+It effectively mirrors stat(2) returned data.
+
+
+@ext - path length to the object.
+@size - the same plus size of the netfs_inode_info structure.
+@id - local inode number.
+@start - 0.
+
+
+@NETFS_PAGE_CACHE
+Command is only received by clients. It contains information about
+page to be marked as not up-to-date.
+
+@id - client's inode number.
+@start - last byte of the page to be invalidated. If it is not equal to
+	current inode size, it will be vmtruncated().
+@size - 0
+@ext - 0
+
+
+@NETFS_READ_PAGES
+Used to read multiple contiguous pages in one go.
+
+@start - first byte of the contiguous region to read.
+@size - contains of two fields: lower 8 bits are used to represent page cache shift
+	used by client, another 3 bytes are used to get number of pages.
+@id - local inode number.
+@ext - path length to the object.
+
+
+@NETFS_RENAME
+Used to rename object.
+Attached data is formed into following string: "old_path|new_path".
+
+@id - local inode number.
+@start - parent inode number.
+@size - length of the above string.
+@ext - length of the old path part.
+
+
+@NETFS_CAPABILITIES
+Used to exchange crypto capabilities with server.
+If crypto capabilities are not supported by server, then client will disable it
+or fail (if 'crypto_fail_unsupported' mount options was specified).
+
+@id - superblock index. Used to specify crypto information for group of servers.
+@size - size of the attached capabilities structure.
+@start - 0.
+@size - 0.
+@scsize - 0.
+
+@NETFS_LOCK
+Used to send lock request/release messages. Although it sends byte range request
+and is capable of flushing pages based on that, it is not used, since all Linux
+filesystems lock the whole inode.
+
+@id - lock generation number.
+@start - start of the locked range.
+@size - size of the locked range.
+@ext - lock type: read/write. Not used actually. 15'th bit is used to determine,
+	if it is lock request (1) or release (0).
+
+@NETFS_XATTR_SET
+@NETFS_XATTR_GET
+Used to set/get extended attributes for given inode.
+@id - attribute generation number or xattr setting type
+@start - size of the attribute (request or attached)
+@size - name length, path len and data size for given attribute
+@ext - path length for given object
diff --git a/Documentation/filesystems/porting b/Documentation/filesystems/porting
new file mode 100644
index 000000000..041b0ded8
--- /dev/null
+++ b/Documentation/filesystems/porting
@@ -0,0 +1,636 @@
+Changes since 2.5.0:
+
+---
+[recommended]
+
+New helpers: sb_bread(), sb_getblk(), sb_find_get_block(), set_bh(),
+	sb_set_blocksize() and sb_min_blocksize().
+
+Use them.
+
+(sb_find_get_block() replaces 2.4's get_hash_table())
+
+---
+[recommended]
+
+New methods: ->alloc_inode() and ->destroy_inode().
+
+Remove inode->u.foo_inode_i
+Declare
+	struct foo_inode_info {
+		/* fs-private stuff */
+		struct inode vfs_inode;
+	};
+	static inline struct foo_inode_info *FOO_I(struct inode *inode)
+	{
+		return list_entry(inode, struct foo_inode_info, vfs_inode);
+	}
+
+Use FOO_I(inode) instead of &inode->u.foo_inode_i;
+
+Add foo_alloc_inode() and foo_destroy_inode() - the former should allocate
+foo_inode_info and return the address of ->vfs_inode, the latter should free
+FOO_I(inode) (see in-tree filesystems for examples).
+
+Make them ->alloc_inode and ->destroy_inode in your super_operations.
+
+Keep in mind that now you need explicit initialization of private data
+typically between calling iget_locked() and unlocking the inode.
+
+At some point that will become mandatory.
+
+---
+[mandatory]
+
+Change of file_system_type method (->read_super to ->get_sb)
+
+->read_super() is no more.  Ditto for DECLARE_FSTYPE and DECLARE_FSTYPE_DEV.
+
+Turn your foo_read_super() into a function that would return 0 in case of
+success and negative number in case of error (-EINVAL unless you have more
+informative error value to report).  Call it foo_fill_super().  Now declare
+
+int foo_get_sb(struct file_system_type *fs_type,
+	int flags, const char *dev_name, void *data, struct vfsmount *mnt)
+{
+	return get_sb_bdev(fs_type, flags, dev_name, data, foo_fill_super,
+			   mnt);
+}
+
+(or similar with s/bdev/nodev/ or s/bdev/single/, depending on the kind of
+filesystem).
+
+Replace DECLARE_FSTYPE... with explicit initializer and have ->get_sb set as
+foo_get_sb.
+
+---
+[mandatory]
+
+Locking change: ->s_vfs_rename_sem is taken only by cross-directory renames.
+Most likely there is no need to change anything, but if you relied on
+global exclusion between renames for some internal purpose - you need to
+change your internal locking.  Otherwise exclusion warranties remain the
+same (i.e. parents and victim are locked, etc.).
+
+---
+[informational]
+
+Now we have the exclusion between ->lookup() and directory removal (by
+->rmdir() and ->rename()).  If you used to need that exclusion and do
+it by internal locking (most of filesystems couldn't care less) - you
+can relax your locking.
+
+---
+[mandatory]
+
+->lookup(), ->truncate(), ->create(), ->unlink(), ->mknod(), ->mkdir(),
+->rmdir(), ->link(), ->lseek(), ->symlink(), ->rename()
+and ->readdir() are called without BKL now.  Grab it on entry, drop upon return
+- that will guarantee the same locking you used to have.  If your method or its
+parts do not need BKL - better yet, now you can shift lock_kernel() and
+unlock_kernel() so that they would protect exactly what needs to be
+protected.
+
+---
+[mandatory]
+
+BKL is also moved from around sb operations. BKL should have been shifted into
+individual fs sb_op functions.  If you don't need it, remove it.
+
+---
+[informational]
+
+check for ->link() target not being a directory is done by callers.  Feel
+free to drop it...
+
+---
+[informational]
+
+->link() callers hold ->i_mutex on the object we are linking to.  Some of your
+problems might be over...
+
+---
+[mandatory]
+
+new file_system_type method - kill_sb(superblock).  If you are converting
+an existing filesystem, set it according to ->fs_flags:
+	FS_REQUIRES_DEV		-	kill_block_super
+	FS_LITTER		-	kill_litter_super
+	neither			-	kill_anon_super
+FS_LITTER is gone - just remove it from fs_flags.
+
+---
+[mandatory]
+
+	FS_SINGLE is gone (actually, that had happened back when ->get_sb()
+went in - and hadn't been documented ;-/).  Just remove it from fs_flags
+(and see ->get_sb() entry for other actions).
+
+---
+[mandatory]
+
+->setattr() is called without BKL now.  Caller _always_ holds ->i_mutex, so
+watch for ->i_mutex-grabbing code that might be used by your ->setattr().
+Callers of notify_change() need ->i_mutex now.
+
+---
+[recommended]
+
+New super_block field "struct export_operations *s_export_op" for
+explicit support for exporting, e.g. via NFS.  The structure is fully
+documented at its declaration in include/linux/fs.h, and in
+Documentation/filesystems/nfs/Exporting.
+
+Briefly it allows for the definition of decode_fh and encode_fh operations
+to encode and decode filehandles, and allows the filesystem to use
+a standard helper function for decode_fh, and provide file-system specific
+support for this helper, particularly get_parent.
+
+It is planned that this will be required for exporting once the code
+settles down a bit.
+
+[mandatory]
+
+s_export_op is now required for exporting a filesystem.
+isofs, ext2, ext3, resierfs, fat
+can be used as examples of very different filesystems.
+
+---
+[mandatory]
+
+iget4() and the read_inode2 callback have been superseded by iget5_locked()
+which has the following prototype,
+
+    struct inode *iget5_locked(struct super_block *sb, unsigned long ino,
+				int (*test)(struct inode *, void *),
+				int (*set)(struct inode *, void *),
+				void *data);
+
+'test' is an additional function that can be used when the inode
+number is not sufficient to identify the actual file object. 'set'
+should be a non-blocking function that initializes those parts of a
+newly created inode to allow the test function to succeed. 'data' is
+passed as an opaque value to both test and set functions.
+
+When the inode has been created by iget5_locked(), it will be returned with the
+I_NEW flag set and will still be locked.  The filesystem then needs to finalize
+the initialization. Once the inode is initialized it must be unlocked by
+calling unlock_new_inode().
+
+The filesystem is responsible for setting (and possibly testing) i_ino
+when appropriate. There is also a simpler iget_locked function that
+just takes the superblock and inode number as arguments and does the
+test and set for you.
+
+e.g.
+	inode = iget_locked(sb, ino);
+	if (inode->i_state & I_NEW) {
+		err = read_inode_from_disk(inode);
+		if (err < 0) {
+			iget_failed(inode);
+			return err;
+		}
+		unlock_new_inode(inode);
+	}
+
+Note that if the process of setting up a new inode fails, then iget_failed()
+should be called on the inode to render it dead, and an appropriate error
+should be passed back to the caller.
+
+---
+[recommended]
+
+->getattr() finally getting used.  See instances in nfs, minix, etc.
+
+---
+[mandatory]
+
+->revalidate() is gone.  If your filesystem had it - provide ->getattr()
+and let it call whatever you had as ->revlidate() + (for symlinks that
+had ->revalidate()) add calls in ->follow_link()/->readlink().
+
+---
+[mandatory]
+
+->d_parent changes are not protected by BKL anymore.  Read access is safe
+if at least one of the following is true:
+	* filesystem has no cross-directory rename()
+	* we know that parent had been locked (e.g. we are looking at
+->d_parent of ->lookup() argument).
+	* we are called from ->rename().
+	* the child's ->d_lock is held
+Audit your code and add locking if needed.  Notice that any place that is
+not protected by the conditions above is risky even in the old tree - you
+had been relying on BKL and that's prone to screwups.  Old tree had quite
+a few holes of that kind - unprotected access to ->d_parent leading to
+anything from oops to silent memory corruption.
+
+---
+[mandatory]
+
+	FS_NOMOUNT is gone.  If you use it - just set SB_NOUSER in flags
+(see rootfs for one kind of solution and bdev/socket/pipe for another).
+
+---
+[recommended]
+
+	Use bdev_read_only(bdev) instead of is_read_only(kdev).  The latter
+is still alive, but only because of the mess in drivers/s390/block/dasd.c.
+As soon as it gets fixed is_read_only() will die.
+
+---
+[mandatory]
+
+->permission() is called without BKL now. Grab it on entry, drop upon
+return - that will guarantee the same locking you used to have.  If
+your method or its parts do not need BKL - better yet, now you can
+shift lock_kernel() and unlock_kernel() so that they would protect
+exactly what needs to be protected.
+
+---
+[mandatory]
+
+->statfs() is now called without BKL held.  BKL should have been
+shifted into individual fs sb_op functions where it's not clear that
+it's safe to remove it.  If you don't need it, remove it.
+
+---
+[mandatory]
+
+	is_read_only() is gone; use bdev_read_only() instead.
+
+---
+[mandatory]
+
+	destroy_buffers() is gone; use invalidate_bdev().
+
+---
+[mandatory]
+
+	fsync_dev() is gone; use fsync_bdev().  NOTE: lvm breakage is
+deliberate; as soon as struct block_device * is propagated in a reasonable
+way by that code fixing will become trivial; until then nothing can be
+done.
+
+[mandatory]
+
+	block truncatation on error exit from ->write_begin, and ->direct_IO
+moved from generic methods (block_write_begin, cont_write_begin,
+nobh_write_begin, blockdev_direct_IO*) to callers.  Take a look at
+ext2_write_failed and callers for an example.
+
+[mandatory]
+
+	->truncate is gone.  The whole truncate sequence needs to be
+implemented in ->setattr, which is now mandatory for filesystems
+implementing on-disk size changes.  Start with a copy of the old inode_setattr
+and vmtruncate, and the reorder the vmtruncate + foofs_vmtruncate sequence to
+be in order of zeroing blocks using block_truncate_page or similar helpers,
+size update and on finally on-disk truncation which should not fail.
+setattr_prepare (which used to be inode_change_ok) now includes the size checks
+for ATTR_SIZE and must be called in the beginning of ->setattr unconditionally.
+
+[mandatory]
+
+	->clear_inode() and ->delete_inode() are gone; ->evict_inode() should
+be used instead.  It gets called whenever the inode is evicted, whether it has
+remaining links or not.  Caller does *not* evict the pagecache or inode-associated
+metadata buffers; the method has to use truncate_inode_pages_final() to get rid
+of those. Caller makes sure async writeback cannot be running for the inode while
+(or after) ->evict_inode() is called.
+
+	->drop_inode() returns int now; it's called on final iput() with
+inode->i_lock held and it returns true if filesystems wants the inode to be
+dropped.  As before, generic_drop_inode() is still the default and it's been
+updated appropriately.  generic_delete_inode() is also alive and it consists
+simply of return 1.  Note that all actual eviction work is done by caller after
+->drop_inode() returns.
+
+	As before, clear_inode() must be called exactly once on each call of
+->evict_inode() (as it used to be for each call of ->delete_inode()).  Unlike
+before, if you are using inode-associated metadata buffers (i.e.
+mark_buffer_dirty_inode()), it's your responsibility to call
+invalidate_inode_buffers() before clear_inode().
+
+	NOTE: checking i_nlink in the beginning of ->write_inode() and bailing out
+if it's zero is not *and* *never* *had* *been* enough.  Final unlink() and iput()
+may happen while the inode is in the middle of ->write_inode(); e.g. if you blindly
+free the on-disk inode, you may end up doing that while ->write_inode() is writing
+to it.
+
+---
+[mandatory]
+
+	.d_delete() now only advises the dcache as to whether or not to cache
+unreferenced dentries, and is now only called when the dentry refcount goes to
+0. Even on 0 refcount transition, it must be able to tolerate being called 0,
+1, or more times (eg. constant, idempotent).
+
+---
+[mandatory]
+
+	.d_compare() calling convention and locking rules are significantly
+changed. Read updated documentation in Documentation/filesystems/vfs.txt (and
+look at examples of other filesystems) for guidance.
+
+---
+[mandatory]
+
+	.d_hash() calling convention and locking rules are significantly
+changed. Read updated documentation in Documentation/filesystems/vfs.txt (and
+look at examples of other filesystems) for guidance.
+
+---
+[mandatory]
+	dcache_lock is gone, replaced by fine grained locks. See fs/dcache.c
+for details of what locks to replace dcache_lock with in order to protect
+particular things. Most of the time, a filesystem only needs ->d_lock, which
+protects *all* the dcache state of a given dentry.
+
+--
+[mandatory]
+
+	Filesystems must RCU-free their inodes, if they can have been accessed
+via rcu-walk path walk (basically, if the file can have had a path name in the
+vfs namespace).
+
+	Even though i_dentry and i_rcu share storage in a union, we will
+initialize the former in inode_init_always(), so just leave it alone in
+the callback.  It used to be necessary to clean it there, but not anymore
+(starting at 3.2).
+
+--
+[recommended]
+	vfs now tries to do path walking in "rcu-walk mode", which avoids
+atomic operations and scalability hazards on dentries and inodes (see
+Documentation/filesystems/path-lookup.txt). d_hash and d_compare changes
+(above) are examples of the changes required to support this. For more complex
+filesystem callbacks, the vfs drops out of rcu-walk mode before the fs call, so
+no changes are required to the filesystem. However, this is costly and loses
+the benefits of rcu-walk mode. We will begin to add filesystem callbacks that
+are rcu-walk aware, shown below. Filesystems should take advantage of this
+where possible.
+
+--
+[mandatory]
+	d_revalidate is a callback that is made on every path element (if
+the filesystem provides it), which requires dropping out of rcu-walk mode. This
+may now be called in rcu-walk mode (nd->flags & LOOKUP_RCU). -ECHILD should be
+returned if the filesystem cannot handle rcu-walk. See
+Documentation/filesystems/vfs.txt for more details.
+
+	permission is an inode permission check that is called on many or all
+directory inodes on the way down a path walk (to check for exec permission). It
+must now be rcu-walk aware (mask & MAY_NOT_BLOCK).  See
+Documentation/filesystems/vfs.txt for more details.
+ 
+--
+[mandatory]
+	In ->fallocate() you must check the mode option passed in.  If your
+filesystem does not support hole punching (deallocating space in the middle of a
+file) you must return -EOPNOTSUPP if FALLOC_FL_PUNCH_HOLE is set in mode.
+Currently you can only have FALLOC_FL_PUNCH_HOLE with FALLOC_FL_KEEP_SIZE set,
+so the i_size should not change when hole punching, even when puching the end of
+a file off.
+
+--
+[mandatory]
+	->get_sb() is gone.  Switch to use of ->mount().  Typically it's just
+a matter of switching from calling get_sb_... to mount_... and changing the
+function type.  If you were doing it manually, just switch from setting ->mnt_root
+to some pointer to returning that pointer.  On errors return ERR_PTR(...).
+
+--
+[mandatory]
+	->permission() and generic_permission()have lost flags
+argument; instead of passing IPERM_FLAG_RCU we add MAY_NOT_BLOCK into mask.
+	generic_permission() has also lost the check_acl argument; ACL checking
+has been taken to VFS and filesystems need to provide a non-NULL ->i_op->get_acl
+to read an ACL from disk.
+
+--
+[mandatory]
+	If you implement your own ->llseek() you must handle SEEK_HOLE and
+SEEK_DATA.  You can hanle this by returning -EINVAL, but it would be nicer to
+support it in some way.  The generic handler assumes that the entire file is
+data and there is a virtual hole at the end of the file.  So if the provided
+offset is less than i_size and SEEK_DATA is specified, return the same offset.
+If the above is true for the offset and you are given SEEK_HOLE, return the end
+of the file.  If the offset is i_size or greater return -ENXIO in either case.
+
+[mandatory]
+	If you have your own ->fsync() you must make sure to call
+filemap_write_and_wait_range() so that all dirty pages are synced out properly.
+You must also keep in mind that ->fsync() is not called with i_mutex held
+anymore, so if you require i_mutex locking you must make sure to take it and
+release it yourself.
+
+--
+[mandatory]
+	d_alloc_root() is gone, along with a lot of bugs caused by code
+misusing it.  Replacement: d_make_root(inode).  The difference is,
+d_make_root() drops the reference to inode if dentry allocation fails.  
+
+--
+[mandatory]
+	The witch is dead!  Well, 2/3 of it, anyway.  ->d_revalidate() and
+->lookup() do *not* take struct nameidata anymore; just the flags.
+--
+[mandatory]
+	->create() doesn't take struct nameidata *; unlike the previous
+two, it gets "is it an O_EXCL or equivalent?" boolean argument.  Note that
+local filesystems can ignore tha argument - they are guaranteed that the
+object doesn't exist.  It's remote/distributed ones that might care...
+--
+[mandatory]
+	FS_REVAL_DOT is gone; if you used to have it, add ->d_weak_revalidate()
+in your dentry operations instead.
+--
+[mandatory]
+	vfs_readdir() is gone; switch to iterate_dir() instead
+--
+[mandatory]
+	->readdir() is gone now; switch to ->iterate()
+[mandatory]
+	vfs_follow_link has been removed.  Filesystems must use nd_set_link
+	from ->follow_link for normal symlinks, or nd_jump_link for magic
+	/proc/<pid> style links.
+--
+[mandatory]
+	iget5_locked()/ilookup5()/ilookup5_nowait() test() callback used to be
+	called with both ->i_lock and inode_hash_lock held; the former is *not*
+	taken anymore, so verify that your callbacks do not rely on it (none
+	of the in-tree instances did).  inode_hash_lock is still held,
+	of course, so they are still serialized wrt removal from inode hash,
+	as well as wrt set() callback of iget5_locked().
+--
+[mandatory]
+	d_materialise_unique() is gone; d_splice_alias() does everything you
+	need now.  Remember that they have opposite orders of arguments ;-/
+--
+[mandatory]
+	f_dentry is gone; use f_path.dentry, or, better yet, see if you can avoid
+	it entirely.
+--
+[mandatory]
+	never call ->read() and ->write() directly; use __vfs_{read,write} or
+	wrappers; instead of checking for ->write or ->read being NULL, look for
+	FMODE_CAN_{WRITE,READ} in file->f_mode.
+--
+[mandatory]
+	do _not_ use new_sync_{read,write} for ->read/->write; leave it NULL
+	instead.
+--
+[mandatory]
+	->aio_read/->aio_write are gone.  Use ->read_iter/->write_iter.
+---
+[recommended]
+	for embedded ("fast") symlinks just set inode->i_link to wherever the
+	symlink body is and use simple_follow_link() as ->follow_link().
+--
+[mandatory]
+	calling conventions for ->follow_link() have changed.  Instead of returning
+	cookie and using nd_set_link() to store the body to traverse, we return
+	the body to traverse and store the cookie using explicit void ** argument.
+	nameidata isn't passed at all - nd_jump_link() doesn't need it and
+	nd_[gs]et_link() is gone.
+--
+[mandatory]
+	calling conventions for ->put_link() have changed.  It gets inode instead of
+	dentry,  it does not get nameidata at all and it gets called only when cookie
+	is non-NULL.  Note that link body isn't available anymore, so if you need it,
+	store it as cookie.
+--
+[mandatory]
+	any symlink that might use page_follow_link_light/page_put_link() must
+	have inode_nohighmem(inode) called before anything might start playing with
+	its pagecache.  No highmem pages should end up in the pagecache of such
+	symlinks.  That includes any preseeding that might be done during symlink
+	creation.  __page_symlink() will honour the mapping gfp flags, so once
+	you've done inode_nohighmem() it's safe to use, but if you allocate and
+	insert the page manually, make sure to use the right gfp flags.
+--
+[mandatory]
+	->follow_link() is replaced with ->get_link(); same API, except that
+		* ->get_link() gets inode as a separate argument
+		* ->get_link() may be called in RCU mode - in that case NULL
+		  dentry is passed
+--
+[mandatory]
+	->get_link() gets struct delayed_call *done now, and should do
+	set_delayed_call() where it used to set *cookie.
+	->put_link() is gone - just give the destructor to set_delayed_call()
+	in ->get_link().
+--
+[mandatory]
+	->getxattr() and xattr_handler.get() get dentry and inode passed separately.
+	dentry might be yet to be attached to inode, so do _not_ use its ->d_inode
+	in the instances.  Rationale: !@#!@# security_d_instantiate() needs to be
+	called before we attach dentry to inode.
+--
+[mandatory]
+	symlinks are no longer the only inodes that do *not* have i_bdev/i_cdev/
+	i_pipe/i_link union zeroed out at inode eviction.  As the result, you can't
+	assume that non-NULL value in ->i_nlink at ->destroy_inode() implies that
+	it's a symlink.  Checking ->i_mode is really needed now.  In-tree we had
+	to fix shmem_destroy_callback() that used to take that kind of shortcut;
+	watch out, since that shortcut is no longer valid.
+--
+[mandatory]
+	->i_mutex is replaced with ->i_rwsem now.  inode_lock() et.al. work as
+	they used to - they just take it exclusive.  However, ->lookup() may be
+	called with parent locked shared.  Its instances must not
+		* use d_instantiate) and d_rehash() separately - use d_add() or
+		  d_splice_alias() instead.
+		* use d_rehash() alone - call d_add(new_dentry, NULL) instead.
+		* in the unlikely case when (read-only) access to filesystem
+		  data structures needs exclusion for some reason, arrange it
+		  yourself.  None of the in-tree filesystems needed that.
+		* rely on ->d_parent and ->d_name not changing after dentry has
+		  been fed to d_add() or d_splice_alias().  Again, none of the
+		  in-tree instances relied upon that.
+	We are guaranteed that lookups of the same name in the same directory
+	will not happen in parallel ("same" in the sense of your ->d_compare()).
+	Lookups on different names in the same directory can and do happen in
+	parallel now.
+--
+[recommended]
+	->iterate_shared() is added; it's a parallel variant of ->iterate().
+	Exclusion on struct file level is still provided (as well as that
+	between it and lseek on the same struct file), but if your directory
+	has been opened several times, you can get these called in parallel.
+	Exclusion between that method and all directory-modifying ones is
+	still provided, of course.
+
+	Often enough ->iterate() can serve as ->iterate_shared() without any
+	changes - it is a read-only operation, after all.  If you have any
+	per-inode or per-dentry in-core data structures modified by ->iterate(),
+	you might need something to serialize the access to them.  If you
+	do dcache pre-seeding, you'll need to switch to d_alloc_parallel() for
+	that; look for in-tree examples.
+
+	Old method is only used if the new one is absent; eventually it will
+	be removed.  Switch while you still can; the old one won't stay.
+--
+[mandatory]
+	->atomic_open() calls without O_CREAT may happen in parallel.
+--
+[mandatory]
+	->setxattr() and xattr_handler.set() get dentry and inode passed separately.
+	dentry might be yet to be attached to inode, so do _not_ use its ->d_inode
+	in the instances.  Rationale: !@#!@# security_d_instantiate() needs to be
+	called before we attach dentry to inode and !@#!@##!@$!$#!@#$!@$!@$ smack
+	->d_instantiate() uses not just ->getxattr() but ->setxattr() as well.
+--
+[mandatory]
+	->d_compare() doesn't get parent as a separate argument anymore.  If you
+	used it for finding the struct super_block involved, dentry->d_sb will
+	work just as well; if it's something more complicated, use dentry->d_parent.
+	Just be careful not to assume that fetching it more than once will yield
+	the same value - in RCU mode it could change under you.
+--
+[mandatory]
+	->rename() has an added flags argument.  Any flags not handled by the
+        filesystem should result in EINVAL being returned.
+--
+[recommended]
+	->readlink is optional for symlinks.  Don't set, unless filesystem needs
+	to fake something for readlink(2).
+--
+[mandatory]
+	->getattr() is now passed a struct path rather than a vfsmount and
+	dentry separately, and it now has request_mask and query_flags arguments
+	to specify the fields and sync type requested by statx.  Filesystems not
+	supporting any statx-specific features may ignore the new arguments.
+--
+[mandatory]
+	->atomic_open() calling conventions have changed.  Gone is int *opened,
+	along with FILE_OPENED/FILE_CREATED.  In place of those we have
+	FMODE_OPENED/FMODE_CREATED, set in file->f_mode.  Additionally, return
+	value for 'called finish_no_open(), open it yourself' case has become
+	0, not 1.  Since finish_no_open() itself is returning 0 now, that part
+	does not need any changes in ->atomic_open() instances.
+--
+[mandatory]
+	alloc_file() has become static now; two wrappers are to be used instead.
+	alloc_file_pseudo(inode, vfsmount, name, flags, ops) is for the cases
+	when dentry needs to be created; that's the majority of old alloc_file()
+	users.  Calling conventions: on success a reference to new struct file
+	is returned and callers reference to inode is subsumed by that.  On
+	failure, ERR_PTR() is returned and no caller's references are affected,
+	so the caller needs to drop the inode reference it held.
+	alloc_file_clone(file, flags, ops) does not affect any caller's references.
+	On success you get a new struct file sharing the mount/dentry with the
+	original, on failure - ERR_PTR().
+--
+[mandatory]
+	DCACHE_RCUACCESS is gone; having an RCU delay on dentry freeing is the
+	default.  DCACHE_NORCU opts out, and only d_alloc_pseudo() has any
+	business doing so.
+--
+[mandatory]
+
+	[should've been added in 2016] stale comment in finish_open()
+	nonwithstanding, failure exits in ->atomic_open() instances should
+	*NOT* fput() the file, no matter what.  Everything is handled by the
+	caller.
diff --git a/Documentation/filesystems/proc.txt b/Documentation/filesystems/proc.txt
new file mode 100644
index 000000000..0d0ecc7df
--- /dev/null
+++ b/Documentation/filesystems/proc.txt
@@ -0,0 +1,1968 @@
+------------------------------------------------------------------------------
+                       T H E  /proc   F I L E S Y S T E M
+------------------------------------------------------------------------------
+/proc/sys         Terrehon Bowden <terrehon@pacbell.net>        October 7 1999
+                  Bodo Bauer <bb@ricochet.net>
+
+2.4.x update	  Jorge Nerin <comandante@zaralinux.com>      November 14 2000
+move /proc/sys	  Shen Feng <shen@cn.fujitsu.com>		  April 1 2009
+------------------------------------------------------------------------------
+Version 1.3                                              Kernel version 2.2.12
+					      Kernel version 2.4.0-test11-pre4
+------------------------------------------------------------------------------
+fixes/update part 1.1  Stefani Seibold <stefani@seibold.net>       June 9 2009
+
+Table of Contents
+-----------------
+
+  0     Preface
+  0.1	Introduction/Credits
+  0.2	Legal Stuff
+
+  1	Collecting System Information
+  1.1	Process-Specific Subdirectories
+  1.2	Kernel data
+  1.3	IDE devices in /proc/ide
+  1.4	Networking info in /proc/net
+  1.5	SCSI info
+  1.6	Parallel port info in /proc/parport
+  1.7	TTY info in /proc/tty
+  1.8	Miscellaneous kernel statistics in /proc/stat
+  1.9	Ext4 file system parameters
+
+  2	Modifying System Parameters
+
+  3	Per-Process Parameters
+  3.1	/proc/<pid>/oom_adj & /proc/<pid>/oom_score_adj - Adjust the oom-killer
+								score
+  3.2	/proc/<pid>/oom_score - Display current oom-killer score
+  3.3	/proc/<pid>/io - Display the IO accounting fields
+  3.4	/proc/<pid>/coredump_filter - Core dump filtering settings
+  3.5	/proc/<pid>/mountinfo - Information about mounts
+  3.6	/proc/<pid>/comm  & /proc/<pid>/task/<tid>/comm
+  3.7   /proc/<pid>/task/<tid>/children - Information about task children
+  3.8   /proc/<pid>/fdinfo/<fd> - Information about opened file
+  3.9   /proc/<pid>/map_files - Information about memory mapped files
+  3.10  /proc/<pid>/timerslack_ns - Task timerslack value
+  3.11	/proc/<pid>/patch_state - Livepatch patch operation state
+
+  4	Configuring procfs
+  4.1	Mount options
+
+------------------------------------------------------------------------------
+Preface
+------------------------------------------------------------------------------
+
+0.1 Introduction/Credits
+------------------------
+
+This documentation is  part of a soon (or  so we hope) to be  released book on
+the SuSE  Linux distribution. As  there is  no complete documentation  for the
+/proc file system and we've used  many freely available sources to write these
+chapters, it  seems only fair  to give the work  back to the  Linux community.
+This work is  based on the 2.2.*  kernel version and the  upcoming 2.4.*. I'm
+afraid it's still far from complete, but we  hope it will be useful. As far as
+we know, it is the first 'all-in-one' document about the /proc file system. It
+is focused  on the Intel  x86 hardware,  so if you  are looking for  PPC, ARM,
+SPARC, AXP, etc., features, you probably  won't find what you are looking for.
+It also only covers IPv4 networking, not IPv6 nor other protocols - sorry. But
+additions and patches  are welcome and will  be added to this  document if you
+mail them to Bodo.
+
+We'd like  to  thank Alan Cox, Rik van Riel, and Alexey Kuznetsov and a lot of
+other people for help compiling this documentation. We'd also like to extend a
+special thank  you to Andi Kleen for documentation, which we relied on heavily
+to create  this  document,  as well as the additional information he provided.
+Thanks to  everybody  else  who contributed source or docs to the Linux kernel
+and helped create a great piece of software... :)
+
+If you  have  any comments, corrections or additions, please don't hesitate to
+contact Bodo  Bauer  at  bb@ricochet.net.  We'll  be happy to add them to this
+document.
+
+The   latest   version    of   this   document   is    available   online   at
+http://tldp.org/LDP/Linux-Filesystem-Hierarchy/html/proc.html
+
+If  the above  direction does  not works  for you,  you could  try the  kernel
+mailing  list  at  linux-kernel@vger.kernel.org  and/or try  to  reach  me  at
+comandante@zaralinux.com.
+
+0.2 Legal Stuff
+---------------
+
+We don't  guarantee  the  correctness  of this document, and if you come to us
+complaining about  how  you  screwed  up  your  system  because  of  incorrect
+documentation, we won't feel responsible...
+
+------------------------------------------------------------------------------
+CHAPTER 1: COLLECTING SYSTEM INFORMATION
+------------------------------------------------------------------------------
+
+------------------------------------------------------------------------------
+In This Chapter
+------------------------------------------------------------------------------
+* Investigating  the  properties  of  the  pseudo  file  system  /proc and its
+  ability to provide information on the running Linux system
+* Examining /proc's structure
+* Uncovering  various  information  about the kernel and the processes running
+  on the system
+------------------------------------------------------------------------------
+
+
+The proc  file  system acts as an interface to internal data structures in the
+kernel. It  can  be  used to obtain information about the system and to change
+certain kernel parameters at runtime (sysctl).
+
+First, we'll  take  a  look  at the read-only parts of /proc. In Chapter 2, we
+show you how you can use /proc/sys to change settings.
+
+1.1 Process-Specific Subdirectories
+-----------------------------------
+
+The directory  /proc  contains  (among other things) one subdirectory for each
+process running on the system, which is named after the process ID (PID).
+
+The link  self  points  to  the  process reading the file system. Each process
+subdirectory has the entries listed in Table 1-1.
+
+
+Table 1-1: Process specific entries in /proc
+..............................................................................
+ File		Content
+ clear_refs	Clears page referenced bits shown in smaps output
+ cmdline	Command line arguments
+ cpu		Current and last cpu in which it was executed	(2.4)(smp)
+ cwd		Link to the current working directory
+ environ	Values of environment variables
+ exe		Link to the executable of this process
+ fd		Directory, which contains all file descriptors
+ maps		Memory maps to executables and library files	(2.4)
+ mem		Memory held by this process
+ root		Link to the root directory of this process
+ stat		Process status
+ statm		Process memory status information
+ status		Process status in human readable form
+ wchan		Present with CONFIG_KALLSYMS=y: it shows the kernel function
+		symbol the task is blocked in - or "0" if not blocked.
+ pagemap	Page table
+ stack		Report full stack trace, enable via CONFIG_STACKTRACE
+ smaps		an extension based on maps, showing the memory consumption of
+		each mapping and flags associated with it
+ numa_maps	an extension based on maps, showing the memory locality and
+		binding policy as well as mem usage (in pages) of each mapping.
+..............................................................................
+
+For example, to get the status information of a process, all you have to do is
+read the file /proc/PID/status:
+
+  >cat /proc/self/status
+  Name:   cat
+  State:  R (running)
+  Tgid:   5452
+  Pid:    5452
+  PPid:   743
+  TracerPid:      0						(2.4)
+  Uid:    501     501     501     501
+  Gid:    100     100     100     100
+  FDSize: 256
+  Groups: 100 14 16
+  VmPeak:     5004 kB
+  VmSize:     5004 kB
+  VmLck:         0 kB
+  VmHWM:       476 kB
+  VmRSS:       476 kB
+  RssAnon:             352 kB
+  RssFile:             120 kB
+  RssShmem:              4 kB
+  VmData:      156 kB
+  VmStk:        88 kB
+  VmExe:        68 kB
+  VmLib:      1412 kB
+  VmPTE:        20 kb
+  VmSwap:        0 kB
+  HugetlbPages:          0 kB
+  CoreDumping:    0
+  Threads:        1
+  SigQ:   0/28578
+  SigPnd: 0000000000000000
+  ShdPnd: 0000000000000000
+  SigBlk: 0000000000000000
+  SigIgn: 0000000000000000
+  SigCgt: 0000000000000000
+  CapInh: 00000000fffffeff
+  CapPrm: 0000000000000000
+  CapEff: 0000000000000000
+  CapBnd: ffffffffffffffff
+  NoNewPrivs:     0
+  Seccomp:        0
+  voluntary_ctxt_switches:        0
+  nonvoluntary_ctxt_switches:     1
+
+This shows you nearly the same information you would get if you viewed it with
+the ps  command.  In  fact,  ps  uses  the  proc  file  system  to  obtain its
+information.  But you get a more detailed  view of the  process by reading the
+file /proc/PID/status. It fields are described in table 1-2.
+
+The  statm  file  contains  more  detailed  information about the process
+memory usage. Its seven fields are explained in Table 1-3.  The stat file
+contains details information about the process itself.  Its fields are
+explained in Table 1-4.
+
+(for SMP CONFIG users)
+For making accounting scalable, RSS related information are handled in an
+asynchronous manner and the value may not be very precise. To see a precise
+snapshot of a moment, you can see /proc/<pid>/smaps file and scan page table.
+It's slow but very precise.
+
+Table 1-2: Contents of the status files (as of 4.8)
+..............................................................................
+ Field                       Content
+ Name                        filename of the executable
+ Umask                       file mode creation mask
+ State                       state (R is running, S is sleeping, D is sleeping
+                             in an uninterruptible wait, Z is zombie,
+			     T is traced or stopped)
+ Tgid                        thread group ID
+ Ngid                        NUMA group ID (0 if none)
+ Pid                         process id
+ PPid                        process id of the parent process
+ TracerPid                   PID of process tracing this process (0 if not)
+ Uid                         Real, effective, saved set, and  file system UIDs
+ Gid                         Real, effective, saved set, and  file system GIDs
+ FDSize                      number of file descriptor slots currently allocated
+ Groups                      supplementary group list
+ NStgid                      descendant namespace thread group ID hierarchy
+ NSpid                       descendant namespace process ID hierarchy
+ NSpgid                      descendant namespace process group ID hierarchy
+ NSsid                       descendant namespace session ID hierarchy
+ VmPeak                      peak virtual memory size
+ VmSize                      total program size
+ VmLck                       locked memory size
+ VmPin                       pinned memory size
+ VmHWM                       peak resident set size ("high water mark")
+ VmRSS                       size of memory portions. It contains the three
+                             following parts (VmRSS = RssAnon + RssFile + RssShmem)
+ RssAnon                     size of resident anonymous memory
+ RssFile                     size of resident file mappings
+ RssShmem                    size of resident shmem memory (includes SysV shm,
+                             mapping of tmpfs and shared anonymous mappings)
+ VmData                      size of private data segments
+ VmStk                       size of stack segments
+ VmExe                       size of text segment
+ VmLib                       size of shared library code
+ VmPTE                       size of page table entries
+ VmSwap                      amount of swap used by anonymous private data
+                             (shmem swap usage is not included)
+ HugetlbPages                size of hugetlb memory portions
+ CoreDumping                 process's memory is currently being dumped
+                             (killing the process may lead to a corrupted core)
+ Threads                     number of threads
+ SigQ                        number of signals queued/max. number for queue
+ SigPnd                      bitmap of pending signals for the thread
+ ShdPnd                      bitmap of shared pending signals for the process
+ SigBlk                      bitmap of blocked signals
+ SigIgn                      bitmap of ignored signals
+ SigCgt                      bitmap of caught signals
+ CapInh                      bitmap of inheritable capabilities
+ CapPrm                      bitmap of permitted capabilities
+ CapEff                      bitmap of effective capabilities
+ CapBnd                      bitmap of capabilities bounding set
+ NoNewPrivs                  no_new_privs, like prctl(PR_GET_NO_NEW_PRIV, ...)
+ Seccomp                     seccomp mode, like prctl(PR_GET_SECCOMP, ...)
+ Cpus_allowed                mask of CPUs on which this process may run
+ Cpus_allowed_list           Same as previous, but in "list format"
+ Mems_allowed                mask of memory nodes allowed to this process
+ Mems_allowed_list           Same as previous, but in "list format"
+ voluntary_ctxt_switches     number of voluntary context switches
+ nonvoluntary_ctxt_switches  number of non voluntary context switches
+..............................................................................
+
+Table 1-3: Contents of the statm files (as of 2.6.8-rc3)
+..............................................................................
+ Field    Content
+ size     total program size (pages)		(same as VmSize in status)
+ resident size of memory portions (pages)	(same as VmRSS in status)
+ shared   number of pages that are shared	(i.e. backed by a file, same
+						as RssFile+RssShmem in status)
+ trs      number of pages that are 'code'	(not including libs; broken,
+							includes data segment)
+ lrs      number of pages of library		(always 0 on 2.6)
+ drs      number of pages of data/stack		(including libs; broken,
+							includes library text)
+ dt       number of dirty pages			(always 0 on 2.6)
+..............................................................................
+
+
+Table 1-4: Contents of the stat files (as of 2.6.30-rc7)
+..............................................................................
+ Field          Content
+  pid           process id
+  tcomm         filename of the executable
+  state         state (R is running, S is sleeping, D is sleeping in an
+                uninterruptible wait, Z is zombie, T is traced or stopped)
+  ppid          process id of the parent process
+  pgrp          pgrp of the process
+  sid           session id
+  tty_nr        tty the process uses
+  tty_pgrp      pgrp of the tty
+  flags         task flags
+  min_flt       number of minor faults
+  cmin_flt      number of minor faults with child's
+  maj_flt       number of major faults
+  cmaj_flt      number of major faults with child's
+  utime         user mode jiffies
+  stime         kernel mode jiffies
+  cutime        user mode jiffies with child's
+  cstime        kernel mode jiffies with child's
+  priority      priority level
+  nice          nice level
+  num_threads   number of threads
+  it_real_value	(obsolete, always 0)
+  start_time    time the process started after system boot
+  vsize         virtual memory size
+  rss           resident set memory size
+  rsslim        current limit in bytes on the rss
+  start_code    address above which program text can run
+  end_code      address below which program text can run
+  start_stack   address of the start of the main process stack
+  esp           current value of ESP
+  eip           current value of EIP
+  pending       bitmap of pending signals
+  blocked       bitmap of blocked signals
+  sigign        bitmap of ignored signals
+  sigcatch      bitmap of caught signals
+  0		(place holder, used to be the wchan address, use /proc/PID/wchan instead)
+  0             (place holder)
+  0             (place holder)
+  exit_signal   signal to send to parent thread on exit
+  task_cpu      which CPU the task is scheduled on
+  rt_priority   realtime priority
+  policy        scheduling policy (man sched_setscheduler)
+  blkio_ticks   time spent waiting for block IO
+  gtime         guest time of the task in jiffies
+  cgtime        guest time of the task children in jiffies
+  start_data    address above which program data+bss is placed
+  end_data      address below which program data+bss is placed
+  start_brk     address above which program heap can be expanded with brk()
+  arg_start     address above which program command line is placed
+  arg_end       address below which program command line is placed
+  env_start     address above which program environment is placed
+  env_end       address below which program environment is placed
+  exit_code     the thread's exit_code in the form reported by the waitpid system call
+..............................................................................
+
+The /proc/PID/maps file containing the currently mapped memory regions and
+their access permissions.
+
+The format is:
+
+address           perms offset  dev   inode      pathname
+
+08048000-08049000 r-xp 00000000 03:00 8312       /opt/test
+08049000-0804a000 rw-p 00001000 03:00 8312       /opt/test
+0804a000-0806b000 rw-p 00000000 00:00 0          [heap]
+a7cb1000-a7cb2000 ---p 00000000 00:00 0
+a7cb2000-a7eb2000 rw-p 00000000 00:00 0
+a7eb2000-a7eb3000 ---p 00000000 00:00 0
+a7eb3000-a7ed5000 rw-p 00000000 00:00 0
+a7ed5000-a8008000 r-xp 00000000 03:00 4222       /lib/libc.so.6
+a8008000-a800a000 r--p 00133000 03:00 4222       /lib/libc.so.6
+a800a000-a800b000 rw-p 00135000 03:00 4222       /lib/libc.so.6
+a800b000-a800e000 rw-p 00000000 00:00 0
+a800e000-a8022000 r-xp 00000000 03:00 14462      /lib/libpthread.so.0
+a8022000-a8023000 r--p 00013000 03:00 14462      /lib/libpthread.so.0
+a8023000-a8024000 rw-p 00014000 03:00 14462      /lib/libpthread.so.0
+a8024000-a8027000 rw-p 00000000 00:00 0
+a8027000-a8043000 r-xp 00000000 03:00 8317       /lib/ld-linux.so.2
+a8043000-a8044000 r--p 0001b000 03:00 8317       /lib/ld-linux.so.2
+a8044000-a8045000 rw-p 0001c000 03:00 8317       /lib/ld-linux.so.2
+aff35000-aff4a000 rw-p 00000000 00:00 0          [stack]
+ffffe000-fffff000 r-xp 00000000 00:00 0          [vdso]
+
+where "address" is the address space in the process that it occupies, "perms"
+is a set of permissions:
+
+ r = read
+ w = write
+ x = execute
+ s = shared
+ p = private (copy on write)
+
+"offset" is the offset into the mapping, "dev" is the device (major:minor), and
+"inode" is the inode  on that device.  0 indicates that  no inode is associated
+with the memory region, as the case would be with BSS (uninitialized data).
+The "pathname" shows the name associated file for this mapping.  If the mapping
+is not associated with a file:
+
+ [heap]                   = the heap of the program
+ [stack]                  = the stack of the main process
+ [vdso]                   = the "virtual dynamic shared object",
+                            the kernel system call handler
+
+ or if empty, the mapping is anonymous.
+
+The /proc/PID/smaps is an extension based on maps, showing the memory
+consumption for each of the process's mappings. For each of mappings there
+is a series of lines such as the following:
+
+08048000-080bc000 r-xp 00000000 03:02 13130      /bin/bash
+Size:               1084 kB
+Rss:                 892 kB
+Pss:                 374 kB
+Shared_Clean:        892 kB
+Shared_Dirty:          0 kB
+Private_Clean:         0 kB
+Private_Dirty:         0 kB
+Referenced:          892 kB
+Anonymous:             0 kB
+LazyFree:              0 kB
+AnonHugePages:         0 kB
+ShmemPmdMapped:        0 kB
+Shared_Hugetlb:        0 kB
+Private_Hugetlb:       0 kB
+Swap:                  0 kB
+SwapPss:               0 kB
+KernelPageSize:        4 kB
+MMUPageSize:           4 kB
+Locked:                0 kB
+THPeligible:           0
+VmFlags: rd ex mr mw me dw
+
+the first of these lines shows the same information as is displayed for the
+mapping in /proc/PID/maps.  The remaining lines show the size of the mapping
+(size), the amount of the mapping that is currently resident in RAM (RSS), the
+process' proportional share of this mapping (PSS), the number of clean and
+dirty private pages in the mapping.
+
+The "proportional set size" (PSS) of a process is the count of pages it has
+in memory, where each page is divided by the number of processes sharing it.
+So if a process has 1000 pages all to itself, and 1000 shared with one other
+process, its PSS will be 1500.
+Note that even a page which is part of a MAP_SHARED mapping, but has only
+a single pte mapped, i.e.  is currently used by only one process, is accounted
+as private and not as shared.
+"Referenced" indicates the amount of memory currently marked as referenced or
+accessed.
+"Anonymous" shows the amount of memory that does not belong to any file.  Even
+a mapping associated with a file may contain anonymous pages: when MAP_PRIVATE
+and a page is modified, the file page is replaced by a private anonymous copy.
+"LazyFree" shows the amount of memory which is marked by madvise(MADV_FREE).
+The memory isn't freed immediately with madvise(). It's freed in memory
+pressure if the memory is clean. Please note that the printed value might
+be lower than the real value due to optimizations used in the current
+implementation. If this is not desirable please file a bug report.
+"AnonHugePages" shows the ammount of memory backed by transparent hugepage.
+"ShmemPmdMapped" shows the ammount of shared (shmem/tmpfs) memory backed by
+huge pages.
+"Shared_Hugetlb" and "Private_Hugetlb" show the ammounts of memory backed by
+hugetlbfs page which is *not* counted in "RSS" or "PSS" field for historical
+reasons. And these are not included in {Shared,Private}_{Clean,Dirty} field.
+"Swap" shows how much would-be-anonymous memory is also used, but out on swap.
+For shmem mappings, "Swap" includes also the size of the mapped (and not
+replaced by copy-on-write) part of the underlying shmem object out on swap.
+"SwapPss" shows proportional swap share of this mapping. Unlike "Swap", this
+does not take into account swapped out page of underlying shmem objects.
+"Locked" indicates whether the mapping is locked in memory or not.
+"THPeligible" indicates whether the mapping is eligible for THP pages - 1 if
+true, 0 otherwise.
+
+"VmFlags" field deserves a separate description. This member represents the kernel
+flags associated with the particular virtual memory area in two letter encoded
+manner. The codes are the following:
+    rd  - readable
+    wr  - writeable
+    ex  - executable
+    sh  - shared
+    mr  - may read
+    mw  - may write
+    me  - may execute
+    ms  - may share
+    gd  - stack segment growns down
+    pf  - pure PFN range
+    dw  - disabled write to the mapped file
+    lo  - pages are locked in memory
+    io  - memory mapped I/O area
+    sr  - sequential read advise provided
+    rr  - random read advise provided
+    dc  - do not copy area on fork
+    de  - do not expand area on remapping
+    ac  - area is accountable
+    nr  - swap space is not reserved for the area
+    ht  - area uses huge tlb pages
+    ar  - architecture specific flag
+    dd  - do not include area into core dump
+    sd  - soft-dirty flag
+    mm  - mixed map area
+    hg  - huge page advise flag
+    nh  - no-huge page advise flag
+    mg  - mergable advise flag
+
+Note that there is no guarantee that every flag and associated mnemonic will
+be present in all further kernel releases. Things get changed, the flags may
+be vanished or the reverse -- new added. Interpretation of their meaning
+might change in future as well. So each consumer of these flags has to
+follow each specific kernel version for the exact semantic.
+
+This file is only present if the CONFIG_MMU kernel configuration option is
+enabled.
+
+Note: reading /proc/PID/maps or /proc/PID/smaps is inherently racy (consistent
+output can be achieved only in the single read call).
+This typically manifests when doing partial reads of these files while the
+memory map is being modified.  Despite the races, we do provide the following
+guarantees:
+
+1) The mapped addresses never go backwards, which implies no two
+   regions will ever overlap.
+2) If there is something at a given vaddr during the entirety of the
+   life of the smaps/maps walk, there will be some output for it.
+
+
+The /proc/PID/clear_refs is used to reset the PG_Referenced and ACCESSED/YOUNG
+bits on both physical and virtual pages associated with a process, and the
+soft-dirty bit on pte (see Documentation/admin-guide/mm/soft-dirty.rst
+for details).
+To clear the bits for all the pages associated with the process
+    > echo 1 > /proc/PID/clear_refs
+
+To clear the bits for the anonymous pages associated with the process
+    > echo 2 > /proc/PID/clear_refs
+
+To clear the bits for the file mapped pages associated with the process
+    > echo 3 > /proc/PID/clear_refs
+
+To clear the soft-dirty bit
+    > echo 4 > /proc/PID/clear_refs
+
+To reset the peak resident set size ("high water mark") to the process's
+current value:
+    > echo 5 > /proc/PID/clear_refs
+
+Any other value written to /proc/PID/clear_refs will have no effect.
+
+The /proc/pid/pagemap gives the PFN, which can be used to find the pageflags
+using /proc/kpageflags and number of times a page is mapped using
+/proc/kpagecount. For detailed explanation, see
+Documentation/admin-guide/mm/pagemap.rst.
+
+The /proc/pid/numa_maps is an extension based on maps, showing the memory
+locality and binding policy, as well as the memory usage (in pages) of
+each mapping. The output follows a general format where mapping details get
+summarized separated by blank spaces, one mapping per each file line:
+
+address   policy    mapping details
+
+00400000 default file=/usr/local/bin/app mapped=1 active=0 N3=1 kernelpagesize_kB=4
+00600000 default file=/usr/local/bin/app anon=1 dirty=1 N3=1 kernelpagesize_kB=4
+3206000000 default file=/lib64/ld-2.12.so mapped=26 mapmax=6 N0=24 N3=2 kernelpagesize_kB=4
+320621f000 default file=/lib64/ld-2.12.so anon=1 dirty=1 N3=1 kernelpagesize_kB=4
+3206220000 default file=/lib64/ld-2.12.so anon=1 dirty=1 N3=1 kernelpagesize_kB=4
+3206221000 default anon=1 dirty=1 N3=1 kernelpagesize_kB=4
+3206800000 default file=/lib64/libc-2.12.so mapped=59 mapmax=21 active=55 N0=41 N3=18 kernelpagesize_kB=4
+320698b000 default file=/lib64/libc-2.12.so
+3206b8a000 default file=/lib64/libc-2.12.so anon=2 dirty=2 N3=2 kernelpagesize_kB=4
+3206b8e000 default file=/lib64/libc-2.12.so anon=1 dirty=1 N3=1 kernelpagesize_kB=4
+3206b8f000 default anon=3 dirty=3 active=1 N3=3 kernelpagesize_kB=4
+7f4dc10a2000 default anon=3 dirty=3 N3=3 kernelpagesize_kB=4
+7f4dc10b4000 default anon=2 dirty=2 active=1 N3=2 kernelpagesize_kB=4
+7f4dc1200000 default file=/anon_hugepage\040(deleted) huge anon=1 dirty=1 N3=1 kernelpagesize_kB=2048
+7fff335f0000 default stack anon=3 dirty=3 N3=3 kernelpagesize_kB=4
+7fff3369d000 default mapped=1 mapmax=35 active=0 N3=1 kernelpagesize_kB=4
+
+Where:
+"address" is the starting address for the mapping;
+"policy" reports the NUMA memory policy set for the mapping (see Documentation/admin-guide/mm/numa_memory_policy.rst);
+"mapping details" summarizes mapping data such as mapping type, page usage counters,
+node locality page counters (N0 == node0, N1 == node1, ...) and the kernel page
+size, in KB, that is backing the mapping up.
+
+1.2 Kernel data
+---------------
+
+Similar to  the  process entries, the kernel data files give information about
+the running kernel. The files used to obtain this information are contained in
+/proc and  are  listed  in Table 1-5. Not all of these will be present in your
+system. It  depends  on the kernel configuration and the loaded modules, which
+files are there, and which are missing.
+
+Table 1-5: Kernel info in /proc
+..............................................................................
+ File        Content                                           
+ apm         Advanced power management info                    
+ buddyinfo   Kernel memory allocator information (see text)	(2.5)
+ bus         Directory containing bus specific information     
+ cmdline     Kernel command line                               
+ cpuinfo     Info about the CPU                                
+ devices     Available devices (block and character)           
+ dma         Used DMS channels                                 
+ filesystems Supported filesystems                             
+ driver	     Various drivers grouped here, currently rtc (2.4)
+ execdomains Execdomains, related to security			(2.4)
+ fb	     Frame Buffer devices				(2.4)
+ fs	     File system parameters, currently nfs/exports	(2.4)
+ ide         Directory containing info about the IDE subsystem 
+ interrupts  Interrupt usage                                   
+ iomem	     Memory map						(2.4)
+ ioports     I/O port usage                                    
+ irq	     Masks for irq to cpu affinity			(2.4)(smp?)
+ isapnp	     ISA PnP (Plug&Play) Info				(2.4)
+ kcore       Kernel core image (can be ELF or A.OUT(deprecated in 2.4))   
+ kmsg        Kernel messages                                   
+ ksyms       Kernel symbol table                               
+ loadavg     Load average of last 1, 5 & 15 minutes                
+ locks       Kernel locks                                      
+ meminfo     Memory info                                       
+ misc        Miscellaneous                                     
+ modules     List of loaded modules                            
+ mounts      Mounted filesystems                               
+ net         Networking info (see text)                        
+ pagetypeinfo Additional page allocator information (see text)  (2.5)
+ partitions  Table of partitions known to the system           
+ pci	     Deprecated info of PCI bus (new way -> /proc/bus/pci/,
+             decoupled by lspci					(2.4)
+ rtc         Real time clock                                   
+ scsi        SCSI info (see text)                              
+ slabinfo    Slab pool info                                    
+ softirqs    softirq usage
+ stat        Overall statistics                                
+ swaps       Swap space utilization                            
+ sys         See chapter 2                                     
+ sysvipc     Info of SysVIPC Resources (msg, sem, shm)		(2.4)
+ tty	     Info of tty drivers
+ uptime      Wall clock since boot, combined idle time of all cpus
+ version     Kernel version                                    
+ video	     bttv info of video resources			(2.4)
+ vmallocinfo Show vmalloced areas
+..............................................................................
+
+You can,  for  example,  check  which interrupts are currently in use and what
+they are used for by looking in the file /proc/interrupts:
+
+  > cat /proc/interrupts 
+             CPU0        
+    0:    8728810          XT-PIC  timer 
+    1:        895          XT-PIC  keyboard 
+    2:          0          XT-PIC  cascade 
+    3:     531695          XT-PIC  aha152x 
+    4:    2014133          XT-PIC  serial 
+    5:      44401          XT-PIC  pcnet_cs 
+    8:          2          XT-PIC  rtc 
+   11:          8          XT-PIC  i82365 
+   12:     182918          XT-PIC  PS/2 Mouse 
+   13:          1          XT-PIC  fpu 
+   14:    1232265          XT-PIC  ide0 
+   15:          7          XT-PIC  ide1 
+  NMI:          0 
+
+In 2.4.* a couple of lines where added to this file LOC & ERR (this time is the
+output of a SMP machine):
+
+  > cat /proc/interrupts 
+
+             CPU0       CPU1       
+    0:    1243498    1214548    IO-APIC-edge  timer
+    1:       8949       8958    IO-APIC-edge  keyboard
+    2:          0          0          XT-PIC  cascade
+    5:      11286      10161    IO-APIC-edge  soundblaster
+    8:          1          0    IO-APIC-edge  rtc
+    9:      27422      27407    IO-APIC-edge  3c503
+   12:     113645     113873    IO-APIC-edge  PS/2 Mouse
+   13:          0          0          XT-PIC  fpu
+   14:      22491      24012    IO-APIC-edge  ide0
+   15:       2183       2415    IO-APIC-edge  ide1
+   17:      30564      30414   IO-APIC-level  eth0
+   18:        177        164   IO-APIC-level  bttv
+  NMI:    2457961    2457959 
+  LOC:    2457882    2457881 
+  ERR:       2155
+
+NMI is incremented in this case because every timer interrupt generates a NMI
+(Non Maskable Interrupt) which is used by the NMI Watchdog to detect lockups.
+
+LOC is the local interrupt counter of the internal APIC of every CPU.
+
+ERR is incremented in the case of errors in the IO-APIC bus (the bus that
+connects the CPUs in a SMP system. This means that an error has been detected,
+the IO-APIC automatically retry the transmission, so it should not be a big
+problem, but you should read the SMP-FAQ.
+
+In 2.6.2* /proc/interrupts was expanded again.  This time the goal was for
+/proc/interrupts to display every IRQ vector in use by the system, not
+just those considered 'most important'.  The new vectors are:
+
+  THR -- interrupt raised when a machine check threshold counter
+  (typically counting ECC corrected errors of memory or cache) exceeds
+  a configurable threshold.  Only available on some systems.
+
+  TRM -- a thermal event interrupt occurs when a temperature threshold
+  has been exceeded for the CPU.  This interrupt may also be generated
+  when the temperature drops back to normal.
+
+  SPU -- a spurious interrupt is some interrupt that was raised then lowered
+  by some IO device before it could be fully processed by the APIC.  Hence
+  the APIC sees the interrupt but does not know what device it came from.
+  For this case the APIC will generate the interrupt with a IRQ vector
+  of 0xff. This might also be generated by chipset bugs.
+
+  RES, CAL, TLB -- rescheduling, call and TLB flush interrupts are
+  sent from one CPU to another per the needs of the OS.  Typically,
+  their statistics are used by kernel developers and interested users to
+  determine the occurrence of interrupts of the given type.
+
+The above IRQ vectors are displayed only when relevant.  For example,
+the threshold vector does not exist on x86_64 platforms.  Others are
+suppressed when the system is a uniprocessor.  As of this writing, only
+i386 and x86_64 platforms support the new IRQ vector displays.
+
+Of some interest is the introduction of the /proc/irq directory to 2.4.
+It could be used to set IRQ to CPU affinity, this means that you can "hook" an
+IRQ to only one CPU, or to exclude a CPU of handling IRQs. The contents of the
+irq subdir is one subdir for each IRQ, and two files; default_smp_affinity and
+prof_cpu_mask.
+
+For example 
+  > ls /proc/irq/
+  0  10  12  14  16  18  2  4  6  8  prof_cpu_mask
+  1  11  13  15  17  19  3  5  7  9  default_smp_affinity
+  > ls /proc/irq/0/
+  smp_affinity
+
+smp_affinity is a bitmask, in which you can specify which CPUs can handle the
+IRQ, you can set it by doing:
+
+  > echo 1 > /proc/irq/10/smp_affinity
+
+This means that only the first CPU will handle the IRQ, but you can also echo
+5 which means that only the first and third CPU can handle the IRQ.
+
+The contents of each smp_affinity file is the same by default:
+
+  > cat /proc/irq/0/smp_affinity
+  ffffffff
+
+There is an alternate interface, smp_affinity_list which allows specifying
+a cpu range instead of a bitmask:
+
+  > cat /proc/irq/0/smp_affinity_list
+  1024-1031
+
+The default_smp_affinity mask applies to all non-active IRQs, which are the
+IRQs which have not yet been allocated/activated, and hence which lack a
+/proc/irq/[0-9]* directory.
+
+The node file on an SMP system shows the node to which the device using the IRQ
+reports itself as being attached. This hardware locality information does not
+include information about any possible driver locality preference.
+
+prof_cpu_mask specifies which CPUs are to be profiled by the system wide
+profiler. Default value is ffffffff (all cpus if there are only 32 of them).
+
+The way IRQs are routed is handled by the IO-APIC, and it's Round Robin
+between all the CPUs which are allowed to handle it. As usual the kernel has
+more info than you and does a better job than you, so the defaults are the
+best choice for almost everyone.  [Note this applies only to those IO-APIC's
+that support "Round Robin" interrupt distribution.]
+
+There are  three  more  important subdirectories in /proc: net, scsi, and sys.
+The general  rule  is  that  the  contents,  or  even  the  existence of these
+directories, depend  on your kernel configuration. If SCSI is not enabled, the
+directory scsi  may  not  exist. The same is true with the net, which is there
+only when networking support is present in the running kernel.
+
+The slabinfo  file  gives  information  about  memory usage at the slab level.
+Linux uses  slab  pools for memory management above page level in version 2.2.
+Commonly used  objects  have  their  own  slab  pool (such as network buffers,
+directory cache, and so on).
+
+..............................................................................
+
+> cat /proc/buddyinfo
+
+Node 0, zone      DMA      0      4      5      4      4      3 ...
+Node 0, zone   Normal      1      0      0      1    101      8 ...
+Node 0, zone  HighMem      2      0      0      1      1      0 ...
+
+External fragmentation is a problem under some workloads, and buddyinfo is a
+useful tool for helping diagnose these problems.  Buddyinfo will give you a 
+clue as to how big an area you can safely allocate, or why a previous
+allocation failed.
+
+Each column represents the number of pages of a certain order which are 
+available.  In this case, there are 0 chunks of 2^0*PAGE_SIZE available in 
+ZONE_DMA, 4 chunks of 2^1*PAGE_SIZE in ZONE_DMA, 101 chunks of 2^4*PAGE_SIZE 
+available in ZONE_NORMAL, etc... 
+
+More information relevant to external fragmentation can be found in
+pagetypeinfo.
+
+> cat /proc/pagetypeinfo
+Page block order: 9
+Pages per block:  512
+
+Free pages count per migrate type at order       0      1      2      3      4      5      6      7      8      9     10
+Node    0, zone      DMA, type    Unmovable      0      0      0      1      1      1      1      1      1      1      0
+Node    0, zone      DMA, type  Reclaimable      0      0      0      0      0      0      0      0      0      0      0
+Node    0, zone      DMA, type      Movable      1      1      2      1      2      1      1      0      1      0      2
+Node    0, zone      DMA, type      Reserve      0      0      0      0      0      0      0      0      0      1      0
+Node    0, zone      DMA, type      Isolate      0      0      0      0      0      0      0      0      0      0      0
+Node    0, zone    DMA32, type    Unmovable    103     54     77      1      1      1     11      8      7      1      9
+Node    0, zone    DMA32, type  Reclaimable      0      0      2      1      0      0      0      0      1      0      0
+Node    0, zone    DMA32, type      Movable    169    152    113     91     77     54     39     13      6      1    452
+Node    0, zone    DMA32, type      Reserve      1      2      2      2      2      0      1      1      1      1      0
+Node    0, zone    DMA32, type      Isolate      0      0      0      0      0      0      0      0      0      0      0
+
+Number of blocks type     Unmovable  Reclaimable      Movable      Reserve      Isolate
+Node 0, zone      DMA            2            0            5            1            0
+Node 0, zone    DMA32           41            6          967            2            0
+
+Fragmentation avoidance in the kernel works by grouping pages of different
+migrate types into the same contiguous regions of memory called page blocks.
+A page block is typically the size of the default hugepage size e.g. 2MB on
+X86-64. By keeping pages grouped based on their ability to move, the kernel
+can reclaim pages within a page block to satisfy a high-order allocation.
+
+The pagetypinfo begins with information on the size of a page block. It
+then gives the same type of information as buddyinfo except broken down
+by migrate-type and finishes with details on how many page blocks of each
+type exist.
+
+If min_free_kbytes has been tuned correctly (recommendations made by hugeadm
+from libhugetlbfs https://github.com/libhugetlbfs/libhugetlbfs/), one can
+make an estimate of the likely number of huge pages that can be allocated
+at a given point in time. All the "Movable" blocks should be allocatable
+unless memory has been mlock()'d. Some of the Reclaimable blocks should
+also be allocatable although a lot of filesystem metadata may have to be
+reclaimed to achieve this.
+
+..............................................................................
+
+meminfo:
+
+Provides information about distribution and utilization of memory.  This
+varies by architecture and compile options.  The following is from a
+16GB PIII, which has highmem enabled.  You may not have all of these fields.
+
+> cat /proc/meminfo
+
+MemTotal:     16344972 kB
+MemFree:      13634064 kB
+MemAvailable: 14836172 kB
+Buffers:          3656 kB
+Cached:        1195708 kB
+SwapCached:          0 kB
+Active:         891636 kB
+Inactive:      1077224 kB
+HighTotal:    15597528 kB
+HighFree:     13629632 kB
+LowTotal:       747444 kB
+LowFree:          4432 kB
+SwapTotal:           0 kB
+SwapFree:            0 kB
+Dirty:             968 kB
+Writeback:           0 kB
+AnonPages:      861800 kB
+Mapped:         280372 kB
+Shmem:             644 kB
+Slab:           284364 kB
+SReclaimable:   159856 kB
+SUnreclaim:     124508 kB
+PageTables:      24448 kB
+NFS_Unstable:        0 kB
+Bounce:              0 kB
+WritebackTmp:        0 kB
+CommitLimit:   7669796 kB
+Committed_AS:   100056 kB
+VmallocTotal:   112216 kB
+VmallocUsed:       428 kB
+VmallocChunk:   111088 kB
+Percpu:          62080 kB
+HardwareCorrupted:   0 kB
+AnonHugePages:   49152 kB
+ShmemHugePages:      0 kB
+ShmemPmdMapped:      0 kB
+
+
+    MemTotal: Total usable ram (i.e. physical ram minus a few reserved
+              bits and the kernel binary code)
+     MemFree: The sum of LowFree+HighFree
+MemAvailable: An estimate of how much memory is available for starting new
+              applications, without swapping. Calculated from MemFree,
+              SReclaimable, the size of the file LRU lists, and the low
+              watermarks in each zone.
+              The estimate takes into account that the system needs some
+              page cache to function well, and that not all reclaimable
+              slab will be reclaimable, due to items being in use. The
+              impact of those factors will vary from system to system.
+     Buffers: Relatively temporary storage for raw disk blocks
+              shouldn't get tremendously large (20MB or so)
+      Cached: in-memory cache for files read from the disk (the
+              pagecache).  Doesn't include SwapCached
+  SwapCached: Memory that once was swapped out, is swapped back in but
+              still also is in the swapfile (if memory is needed it
+              doesn't need to be swapped out AGAIN because it is already
+              in the swapfile. This saves I/O)
+      Active: Memory that has been used more recently and usually not
+              reclaimed unless absolutely necessary.
+    Inactive: Memory which has been less recently used.  It is more
+              eligible to be reclaimed for other purposes
+   HighTotal:
+    HighFree: Highmem is all memory above ~860MB of physical memory
+              Highmem areas are for use by userspace programs, or
+              for the pagecache.  The kernel must use tricks to access
+              this memory, making it slower to access than lowmem.
+    LowTotal:
+     LowFree: Lowmem is memory which can be used for everything that
+              highmem can be used for, but it is also available for the
+              kernel's use for its own data structures.  Among many
+              other things, it is where everything from the Slab is
+              allocated.  Bad things happen when you're out of lowmem.
+   SwapTotal: total amount of swap space available
+    SwapFree: Memory which has been evicted from RAM, and is temporarily
+              on the disk
+       Dirty: Memory which is waiting to get written back to the disk
+   Writeback: Memory which is actively being written back to the disk
+   AnonPages: Non-file backed pages mapped into userspace page tables
+HardwareCorrupted: The amount of RAM/memory in KB, the kernel identifies as
+	      corrupted.
+AnonHugePages: Non-file backed huge pages mapped into userspace page tables
+      Mapped: files which have been mmaped, such as libraries
+       Shmem: Total memory used by shared memory (shmem) and tmpfs
+ShmemHugePages: Memory used by shared memory (shmem) and tmpfs allocated
+              with huge pages
+ShmemPmdMapped: Shared memory mapped into userspace with huge pages
+        Slab: in-kernel data structures cache
+SReclaimable: Part of Slab, that might be reclaimed, such as caches
+  SUnreclaim: Part of Slab, that cannot be reclaimed on memory pressure
+  PageTables: amount of memory dedicated to the lowest level of page
+              tables.
+NFS_Unstable: NFS pages sent to the server, but not yet committed to stable
+	      storage
+      Bounce: Memory used for block device "bounce buffers"
+WritebackTmp: Memory used by FUSE for temporary writeback buffers
+ CommitLimit: Based on the overcommit ratio ('vm.overcommit_ratio'),
+              this is the total amount of  memory currently available to
+              be allocated on the system. This limit is only adhered to
+              if strict overcommit accounting is enabled (mode 2 in
+              'vm.overcommit_memory').
+              The CommitLimit is calculated with the following formula:
+              CommitLimit = ([total RAM pages] - [total huge TLB pages]) *
+                             overcommit_ratio / 100 + [total swap pages]
+              For example, on a system with 1G of physical RAM and 7G
+              of swap with a `vm.overcommit_ratio` of 30 it would
+              yield a CommitLimit of 7.3G.
+              For more details, see the memory overcommit documentation
+              in vm/overcommit-accounting.
+Committed_AS: The amount of memory presently allocated on the system.
+              The committed memory is a sum of all of the memory which
+              has been allocated by processes, even if it has not been
+              "used" by them as of yet. A process which malloc()'s 1G
+              of memory, but only touches 300M of it will show up as
+	      using 1G. This 1G is memory which has been "committed" to
+              by the VM and can be used at any time by the allocating
+              application. With strict overcommit enabled on the system
+              (mode 2 in 'vm.overcommit_memory'),allocations which would
+              exceed the CommitLimit (detailed above) will not be permitted.
+              This is useful if one needs to guarantee that processes will
+              not fail due to lack of memory once that memory has been
+              successfully allocated.
+VmallocTotal: total size of vmalloc memory area
+ VmallocUsed: amount of vmalloc area which is used
+VmallocChunk: largest contiguous block of vmalloc area which is free
+      Percpu: Memory allocated to the percpu allocator used to back percpu
+              allocations. This stat excludes the cost of metadata.
+
+..............................................................................
+
+vmallocinfo:
+
+Provides information about vmalloced/vmaped areas. One line per area,
+containing the virtual address range of the area, size in bytes,
+caller information of the creator, and optional information depending
+on the kind of area :
+
+ pages=nr    number of pages
+ phys=addr   if a physical address was specified
+ ioremap     I/O mapping (ioremap() and friends)
+ vmalloc     vmalloc() area
+ vmap        vmap()ed pages
+ user        VM_USERMAP area
+ vpages      buffer for pages pointers was vmalloced (huge area)
+ N<node>=nr  (Only on NUMA kernels)
+             Number of pages allocated on memory node <node>
+
+> cat /proc/vmallocinfo
+0xffffc20000000000-0xffffc20000201000 2101248 alloc_large_system_hash+0x204 ...
+  /0x2c0 pages=512 vmalloc N0=128 N1=128 N2=128 N3=128
+0xffffc20000201000-0xffffc20000302000 1052672 alloc_large_system_hash+0x204 ...
+  /0x2c0 pages=256 vmalloc N0=64 N1=64 N2=64 N3=64
+0xffffc20000302000-0xffffc20000304000    8192 acpi_tb_verify_table+0x21/0x4f...
+  phys=7fee8000 ioremap
+0xffffc20000304000-0xffffc20000307000   12288 acpi_tb_verify_table+0x21/0x4f...
+  phys=7fee7000 ioremap
+0xffffc2000031d000-0xffffc2000031f000    8192 init_vdso_vars+0x112/0x210
+0xffffc2000031f000-0xffffc2000032b000   49152 cramfs_uncompress_init+0x2e ...
+  /0x80 pages=11 vmalloc N0=3 N1=3 N2=2 N3=3
+0xffffc2000033a000-0xffffc2000033d000   12288 sys_swapon+0x640/0xac0      ...
+  pages=2 vmalloc N1=2
+0xffffc20000347000-0xffffc2000034c000   20480 xt_alloc_table_info+0xfe ...
+  /0x130 [x_tables] pages=4 vmalloc N0=4
+0xffffffffa0000000-0xffffffffa000f000   61440 sys_init_module+0xc27/0x1d00 ...
+   pages=14 vmalloc N2=14
+0xffffffffa000f000-0xffffffffa0014000   20480 sys_init_module+0xc27/0x1d00 ...
+   pages=4 vmalloc N1=4
+0xffffffffa0014000-0xffffffffa0017000   12288 sys_init_module+0xc27/0x1d00 ...
+   pages=2 vmalloc N1=2
+0xffffffffa0017000-0xffffffffa0022000   45056 sys_init_module+0xc27/0x1d00 ...
+   pages=10 vmalloc N0=10
+
+..............................................................................
+
+softirqs:
+
+Provides counts of softirq handlers serviced since boot time, for each cpu.
+
+> cat /proc/softirqs
+                CPU0       CPU1       CPU2       CPU3
+      HI:          0          0          0          0
+   TIMER:      27166      27120      27097      27034
+  NET_TX:          0          0          0         17
+  NET_RX:         42          0          0         39
+   BLOCK:          0          0        107       1121
+ TASKLET:          0          0          0        290
+   SCHED:      27035      26983      26971      26746
+ HRTIMER:          0          0          0          0
+     RCU:       1678       1769       2178       2250
+
+
+1.3 IDE devices in /proc/ide
+----------------------------
+
+The subdirectory /proc/ide contains information about all IDE devices of which
+the kernel  is  aware.  There is one subdirectory for each IDE controller, the
+file drivers  and a link for each IDE device, pointing to the device directory
+in the controller specific subtree.
+
+The file  drivers  contains general information about the drivers used for the
+IDE devices:
+
+  > cat /proc/ide/drivers
+  ide-cdrom version 4.53
+  ide-disk version 1.08
+
+More detailed  information  can  be  found  in  the  controller  specific
+subdirectories. These  are  named  ide0,  ide1  and  so  on.  Each  of  these
+directories contains the files shown in table 1-6.
+
+
+Table 1-6: IDE controller info in  /proc/ide/ide?
+..............................................................................
+ File    Content                                 
+ channel IDE channel (0 or 1)                    
+ config  Configuration (only for PCI/IDE bridge) 
+ mate    Mate name                               
+ model   Type/Chipset of IDE controller          
+..............................................................................
+
+Each device  connected  to  a  controller  has  a separate subdirectory in the
+controllers directory.  The  files  listed in table 1-7 are contained in these
+directories.
+
+
+Table 1-7: IDE device information
+..............................................................................
+ File             Content                                    
+ cache            The cache                                  
+ capacity         Capacity of the medium (in 512Byte blocks) 
+ driver           driver and version                         
+ geometry         physical and logical geometry              
+ identify         device identify block                      
+ media            media type                                 
+ model            device identifier                          
+ settings         device setup                               
+ smart_thresholds IDE disk management thresholds             
+ smart_values     IDE disk management values                 
+..............................................................................
+
+The most  interesting  file is settings. This file contains a nice overview of
+the drive parameters:
+
+  # cat /proc/ide/ide0/hda/settings 
+  name                    value           min             max             mode 
+  ----                    -----           ---             ---             ---- 
+  bios_cyl                526             0               65535           rw 
+  bios_head               255             0               255             rw 
+  bios_sect               63              0               63              rw 
+  breada_readahead        4               0               127             rw 
+  bswap                   0               0               1               r 
+  file_readahead          72              0               2097151         rw 
+  io_32bit                0               0               3               rw 
+  keepsettings            0               0               1               rw 
+  max_kb_per_request      122             1               127             rw 
+  multcount               0               0               8               rw 
+  nice1                   1               0               1               rw 
+  nowerr                  0               0               1               rw 
+  pio_mode                write-only      0               255             w 
+  slow                    0               0               1               rw 
+  unmaskirq               0               0               1               rw 
+  using_dma               0               0               1               rw 
+
+
+1.4 Networking info in /proc/net
+--------------------------------
+
+The subdirectory  /proc/net  follows  the  usual  pattern. Table 1-8 shows the
+additional values  you  get  for  IP  version 6 if you configure the kernel to
+support this. Table 1-9 lists the files and their meaning.
+
+
+Table 1-8: IPv6 info in /proc/net
+..............................................................................
+ File       Content                                               
+ udp6       UDP sockets (IPv6)                                    
+ tcp6       TCP sockets (IPv6)                                    
+ raw6       Raw device statistics (IPv6)                          
+ igmp6      IP multicast addresses, which this host joined (IPv6) 
+ if_inet6   List of IPv6 interface addresses                      
+ ipv6_route Kernel routing table for IPv6                         
+ rt6_stats  Global IPv6 routing tables statistics                 
+ sockstat6  Socket statistics (IPv6)                              
+ snmp6      Snmp data (IPv6)                                      
+..............................................................................
+
+
+Table 1-9: Network info in /proc/net
+..............................................................................
+ File          Content                                                         
+ arp           Kernel  ARP table                                               
+ dev           network devices with statistics                                 
+ dev_mcast     the Layer2 multicast groups a device is listening too
+               (interface index, label, number of references, number of bound
+               addresses). 
+ dev_stat      network device status                                           
+ ip_fwchains   Firewall chain linkage                                          
+ ip_fwnames    Firewall chain names                                            
+ ip_masq       Directory containing the masquerading tables                    
+ ip_masquerade Major masquerading table                                        
+ netstat       Network statistics                                              
+ raw           raw device statistics                                           
+ route         Kernel routing table                                            
+ rpc           Directory containing rpc info                                   
+ rt_cache      Routing cache                                                   
+ snmp          SNMP data                                                       
+ sockstat      Socket statistics                                               
+ tcp           TCP  sockets                                                    
+ udp           UDP sockets                                                     
+ unix          UNIX domain sockets                                             
+ wireless      Wireless interface data (Wavelan etc)                           
+ igmp          IP multicast addresses, which this host joined                  
+ psched        Global packet scheduler parameters.                             
+ netlink       List of PF_NETLINK sockets                                      
+ ip_mr_vifs    List of multicast virtual interfaces                            
+ ip_mr_cache   List of multicast routing cache                                 
+..............................................................................
+
+You can  use  this  information  to see which network devices are available in
+your system and how much traffic was routed over those devices:
+
+  > cat /proc/net/dev 
+  Inter-|Receive                                                   |[... 
+   face |bytes    packets errs drop fifo frame compressed multicast|[... 
+      lo:  908188   5596     0    0    0     0          0         0 [...         
+    ppp0:15475140  20721   410    0    0   410          0         0 [...  
+    eth0:  614530   7085     0    0    0     0          0         1 [... 
+   
+  ...] Transmit 
+  ...] bytes    packets errs drop fifo colls carrier compressed 
+  ...]  908188     5596    0    0    0     0       0          0 
+  ...] 1375103    17405    0    0    0     0       0          0 
+  ...] 1703981     5535    0    0    0     3       0          0 
+
+In addition, each Channel Bond interface has its own directory.  For
+example, the bond0 device will have a directory called /proc/net/bond0/.
+It will contain information that is specific to that bond, such as the
+current slaves of the bond, the link status of the slaves, and how
+many times the slaves link has failed.
+
+1.5 SCSI info
+-------------
+
+If you  have  a  SCSI  host adapter in your system, you'll find a subdirectory
+named after  the driver for this adapter in /proc/scsi. You'll also see a list
+of all recognized SCSI devices in /proc/scsi:
+
+  >cat /proc/scsi/scsi 
+  Attached devices: 
+  Host: scsi0 Channel: 00 Id: 00 Lun: 00 
+    Vendor: IBM      Model: DGHS09U          Rev: 03E0 
+    Type:   Direct-Access                    ANSI SCSI revision: 03 
+  Host: scsi0 Channel: 00 Id: 06 Lun: 00 
+    Vendor: PIONEER  Model: CD-ROM DR-U06S   Rev: 1.04 
+    Type:   CD-ROM                           ANSI SCSI revision: 02 
+
+
+The directory  named  after  the driver has one file for each adapter found in
+the system.  These  files  contain information about the controller, including
+the used  IRQ  and  the  IO  address range. The amount of information shown is
+dependent on  the adapter you use. The example shows the output for an Adaptec
+AHA-2940 SCSI adapter:
+
+  > cat /proc/scsi/aic7xxx/0 
+   
+  Adaptec AIC7xxx driver version: 5.1.19/3.2.4 
+  Compile Options: 
+    TCQ Enabled By Default : Disabled 
+    AIC7XXX_PROC_STATS     : Disabled 
+    AIC7XXX_RESET_DELAY    : 5 
+  Adapter Configuration: 
+             SCSI Adapter: Adaptec AHA-294X Ultra SCSI host adapter 
+                             Ultra Wide Controller 
+      PCI MMAPed I/O Base: 0xeb001000 
+   Adapter SEEPROM Config: SEEPROM found and used. 
+        Adaptec SCSI BIOS: Enabled 
+                      IRQ: 10 
+                     SCBs: Active 0, Max Active 2, 
+                           Allocated 15, HW 16, Page 255 
+               Interrupts: 160328 
+        BIOS Control Word: 0x18b6 
+     Adapter Control Word: 0x005b 
+     Extended Translation: Enabled 
+  Disconnect Enable Flags: 0xffff 
+       Ultra Enable Flags: 0x0001 
+   Tag Queue Enable Flags: 0x0000 
+  Ordered Queue Tag Flags: 0x0000 
+  Default Tag Queue Depth: 8 
+      Tagged Queue By Device array for aic7xxx host instance 0: 
+        {255,255,255,255,255,255,255,255,255,255,255,255,255,255,255,255} 
+      Actual queue depth per device for aic7xxx host instance 0: 
+        {1,1,1,1,1,1,1,1,1,1,1,1,1,1,1,1} 
+  Statistics: 
+  (scsi0:0:0:0) 
+    Device using Wide/Sync transfers at 40.0 MByte/sec, offset 8 
+    Transinfo settings: current(12/8/1/0), goal(12/8/1/0), user(12/15/1/0) 
+    Total transfers 160151 (74577 reads and 85574 writes) 
+  (scsi0:0:6:0) 
+    Device using Narrow/Sync transfers at 5.0 MByte/sec, offset 15 
+    Transinfo settings: current(50/15/0/0), goal(50/15/0/0), user(50/15/0/0) 
+    Total transfers 0 (0 reads and 0 writes) 
+
+
+1.6 Parallel port info in /proc/parport
+---------------------------------------
+
+The directory  /proc/parport  contains information about the parallel ports of
+your system.  It  has  one  subdirectory  for  each port, named after the port
+number (0,1,2,...).
+
+These directories contain the four files shown in Table 1-10.
+
+
+Table 1-10: Files in /proc/parport
+..............................................................................
+ File      Content                                                             
+ autoprobe Any IEEE-1284 device ID information that has been acquired.         
+ devices   list of the device drivers using that port. A + will appear by the
+           name of the device currently using the port (it might not appear
+           against any). 
+ hardware  Parallel port's base address, IRQ line and DMA channel.             
+ irq       IRQ that parport is using for that port. This is in a separate
+           file to allow you to alter it by writing a new value in (IRQ
+           number or none). 
+..............................................................................
+
+1.7 TTY info in /proc/tty
+-------------------------
+
+Information about  the  available  and actually used tty's can be found in the
+directory /proc/tty.You'll  find  entries  for drivers and line disciplines in
+this directory, as shown in Table 1-11.
+
+
+Table 1-11: Files in /proc/tty
+..............................................................................
+ File          Content                                        
+ drivers       list of drivers and their usage                
+ ldiscs        registered line disciplines                    
+ driver/serial usage statistic and status of single tty lines 
+..............................................................................
+
+To see  which  tty's  are  currently in use, you can simply look into the file
+/proc/tty/drivers:
+
+  > cat /proc/tty/drivers 
+  pty_slave            /dev/pts      136   0-255 pty:slave 
+  pty_master           /dev/ptm      128   0-255 pty:master 
+  pty_slave            /dev/ttyp       3   0-255 pty:slave 
+  pty_master           /dev/pty        2   0-255 pty:master 
+  serial               /dev/cua        5   64-67 serial:callout 
+  serial               /dev/ttyS       4   64-67 serial 
+  /dev/tty0            /dev/tty0       4       0 system:vtmaster 
+  /dev/ptmx            /dev/ptmx       5       2 system 
+  /dev/console         /dev/console    5       1 system:console 
+  /dev/tty             /dev/tty        5       0 system:/dev/tty 
+  unknown              /dev/tty        4    1-63 console 
+
+
+1.8 Miscellaneous kernel statistics in /proc/stat
+-------------------------------------------------
+
+Various pieces   of  information about  kernel activity  are  available in the
+/proc/stat file.  All  of  the numbers reported  in  this file are  aggregates
+since the system first booted.  For a quick look, simply cat the file:
+
+  > cat /proc/stat
+  cpu  2255 34 2290 22625563 6290 127 456 0 0 0
+  cpu0 1132 34 1441 11311718 3675 127 438 0 0 0
+  cpu1 1123 0 849 11313845 2614 0 18 0 0 0
+  intr 114930548 113199788 3 0 5 263 0 4 [... lots more numbers ...]
+  ctxt 1990473
+  btime 1062191376
+  processes 2915
+  procs_running 1
+  procs_blocked 0
+  softirq 183433 0 21755 12 39 1137 231 21459 2263
+
+The very first  "cpu" line aggregates the  numbers in all  of the other "cpuN"
+lines.  These numbers identify the amount of time the CPU has spent performing
+different kinds of work.  Time units are in USER_HZ (typically hundredths of a
+second).  The meanings of the columns are as follows, from left to right:
+
+- user: normal processes executing in user mode
+- nice: niced processes executing in user mode
+- system: processes executing in kernel mode
+- idle: twiddling thumbs
+- iowait: In a word, iowait stands for waiting for I/O to complete. But there
+  are several problems:
+  1. Cpu will not wait for I/O to complete, iowait is the time that a task is
+     waiting for I/O to complete. When cpu goes into idle state for
+     outstanding task io, another task will be scheduled on this CPU.
+  2. In a multi-core CPU, the task waiting for I/O to complete is not running
+     on any CPU, so the iowait of each CPU is difficult to calculate.
+  3. The value of iowait field in /proc/stat will decrease in certain
+     conditions.
+  So, the iowait is not reliable by reading from /proc/stat.
+- irq: servicing interrupts
+- softirq: servicing softirqs
+- steal: involuntary wait
+- guest: running a normal guest
+- guest_nice: running a niced guest
+
+The "intr" line gives counts of interrupts  serviced since boot time, for each
+of the  possible system interrupts.   The first  column  is the  total of  all
+interrupts serviced  including  unnumbered  architecture specific  interrupts;
+each  subsequent column is the  total for that particular numbered interrupt.
+Unnumbered interrupts are not shown, only summed into the total.
+
+The "ctxt" line gives the total number of context switches across all CPUs.
+
+The "btime" line gives  the time at which the  system booted, in seconds since
+the Unix epoch.
+
+The "processes" line gives the number  of processes and threads created, which
+includes (but  is not limited  to) those  created by  calls to the  fork() and
+clone() system calls.
+
+The "procs_running" line gives the total number of threads that are
+running or ready to run (i.e., the total number of runnable threads).
+
+The   "procs_blocked" line gives  the  number of  processes currently blocked,
+waiting for I/O to complete.
+
+The "softirq" line gives counts of softirqs serviced since boot time, for each
+of the possible system softirqs. The first column is the total of all
+softirqs serviced; each subsequent column is the total for that particular
+softirq.
+
+
+1.9 Ext4 file system parameters
+-------------------------------
+
+Information about mounted ext4 file systems can be found in
+/proc/fs/ext4.  Each mounted filesystem will have a directory in
+/proc/fs/ext4 based on its device name (i.e., /proc/fs/ext4/hdc or
+/proc/fs/ext4/dm-0).   The files in each per-device directory are shown
+in Table 1-12, below.
+
+Table 1-12: Files in /proc/fs/ext4/<devname>
+..............................................................................
+ File            Content                                        
+ mb_groups       details of multiblock allocator buddy cache of free blocks
+..............................................................................
+
+2.0 /proc/consoles
+------------------
+Shows registered system console lines.
+
+To see which character device lines are currently used for the system console
+/dev/console, you may simply look into the file /proc/consoles:
+
+  > cat /proc/consoles
+  tty0                 -WU (ECp)       4:7
+  ttyS0                -W- (Ep)        4:64
+
+The columns are:
+
+  device               name of the device
+  operations           R = can do read operations
+                       W = can do write operations
+                       U = can do unblank
+  flags                E = it is enabled
+                       C = it is preferred console
+                       B = it is primary boot console
+                       p = it is used for printk buffer
+                       b = it is not a TTY but a Braille device
+                       a = it is safe to use when cpu is offline
+  major:minor          major and minor number of the device separated by a colon
+
+------------------------------------------------------------------------------
+Summary
+------------------------------------------------------------------------------
+The /proc file system serves information about the running system. It not only
+allows access to process data but also allows you to request the kernel status
+by reading files in the hierarchy.
+
+The directory  structure  of /proc reflects the types of information and makes
+it easy, if not obvious, where to look for specific data.
+------------------------------------------------------------------------------
+
+------------------------------------------------------------------------------
+CHAPTER 2: MODIFYING SYSTEM PARAMETERS
+------------------------------------------------------------------------------
+
+------------------------------------------------------------------------------
+In This Chapter
+------------------------------------------------------------------------------
+* Modifying kernel parameters by writing into files found in /proc/sys
+* Exploring the files which modify certain parameters
+* Review of the /proc/sys file tree
+------------------------------------------------------------------------------
+
+
+A very  interesting part of /proc is the directory /proc/sys. This is not only
+a source  of  information,  it also allows you to change parameters within the
+kernel. Be  very  careful  when attempting this. You can optimize your system,
+but you  can  also  cause  it  to  crash.  Never  alter kernel parameters on a
+production system.  Set  up  a  development machine and test to make sure that
+everything works  the  way  you want it to. You may have no alternative but to
+reboot the machine once an error has been made.
+
+To change  a  value,  simply  echo  the new value into the file. An example is
+given below  in the section on the file system data. You need to be root to do
+this. You  can  create  your  own  boot script to perform this every time your
+system boots.
+
+The files  in /proc/sys can be used to fine tune and monitor miscellaneous and
+general things  in  the operation of the Linux kernel. Since some of the files
+can inadvertently  disrupt  your  system,  it  is  advisable  to  read  both
+documentation and  source  before actually making adjustments. In any case, be
+very careful  when  writing  to  any  of these files. The entries in /proc may
+change slightly between the 2.1.* and the 2.2 kernel, so if there is any doubt
+review the kernel documentation in the directory /usr/src/linux/Documentation.
+This chapter  is  heavily  based  on the documentation included in the pre 2.2
+kernels, and became part of it in version 2.2.1 of the Linux kernel.
+
+Please see: Documentation/sysctl/ directory for descriptions of these
+entries.
+
+------------------------------------------------------------------------------
+Summary
+------------------------------------------------------------------------------
+Certain aspects  of  kernel  behavior  can be modified at runtime, without the
+need to  recompile  the kernel, or even to reboot the system. The files in the
+/proc/sys tree  can  not only be read, but also modified. You can use the echo
+command to write value into these files, thereby changing the default settings
+of the kernel.
+------------------------------------------------------------------------------
+
+------------------------------------------------------------------------------
+CHAPTER 3: PER-PROCESS PARAMETERS
+------------------------------------------------------------------------------
+
+3.1 /proc/<pid>/oom_adj & /proc/<pid>/oom_score_adj- Adjust the oom-killer score
+--------------------------------------------------------------------------------
+
+These file can be used to adjust the badness heuristic used to select which
+process gets killed in out of memory conditions.
+
+The badness heuristic assigns a value to each candidate task ranging from 0
+(never kill) to 1000 (always kill) to determine which process is targeted.  The
+units are roughly a proportion along that range of allowed memory the process
+may allocate from based on an estimation of its current memory and swap use.
+For example, if a task is using all allowed memory, its badness score will be
+1000.  If it is using half of its allowed memory, its score will be 500.
+
+There is an additional factor included in the badness score: the current memory
+and swap usage is discounted by 3% for root processes.
+
+The amount of "allowed" memory depends on the context in which the oom killer
+was called.  If it is due to the memory assigned to the allocating task's cpuset
+being exhausted, the allowed memory represents the set of mems assigned to that
+cpuset.  If it is due to a mempolicy's node(s) being exhausted, the allowed
+memory represents the set of mempolicy nodes.  If it is due to a memory
+limit (or swap limit) being reached, the allowed memory is that configured
+limit.  Finally, if it is due to the entire system being out of memory, the
+allowed memory represents all allocatable resources.
+
+The value of /proc/<pid>/oom_score_adj is added to the badness score before it
+is used to determine which task to kill.  Acceptable values range from -1000
+(OOM_SCORE_ADJ_MIN) to +1000 (OOM_SCORE_ADJ_MAX).  This allows userspace to
+polarize the preference for oom killing either by always preferring a certain
+task or completely disabling it.  The lowest possible value, -1000, is
+equivalent to disabling oom killing entirely for that task since it will always
+report a badness score of 0.
+
+Consequently, it is very simple for userspace to define the amount of memory to
+consider for each task.  Setting a /proc/<pid>/oom_score_adj value of +500, for
+example, is roughly equivalent to allowing the remainder of tasks sharing the
+same system, cpuset, mempolicy, or memory controller resources to use at least
+50% more memory.  A value of -500, on the other hand, would be roughly
+equivalent to discounting 50% of the task's allowed memory from being considered
+as scoring against the task.
+
+For backwards compatibility with previous kernels, /proc/<pid>/oom_adj may also
+be used to tune the badness score.  Its acceptable values range from -16
+(OOM_ADJUST_MIN) to +15 (OOM_ADJUST_MAX) and a special value of -17
+(OOM_DISABLE) to disable oom killing entirely for that task.  Its value is
+scaled linearly with /proc/<pid>/oom_score_adj.
+
+The value of /proc/<pid>/oom_score_adj may be reduced no lower than the last
+value set by a CAP_SYS_RESOURCE process. To reduce the value any lower
+requires CAP_SYS_RESOURCE.
+
+Caveat: when a parent task is selected, the oom killer will sacrifice any first
+generation children with separate address spaces instead, if possible.  This
+avoids servers and important system daemons from being killed and loses the
+minimal amount of work.
+
+
+3.2 /proc/<pid>/oom_score - Display current oom-killer score
+-------------------------------------------------------------
+
+This file can be used to check the current score used by the oom-killer is for
+any given <pid>. Use it together with /proc/<pid>/oom_score_adj to tune which
+process should be killed in an out-of-memory situation.
+
+
+3.3  /proc/<pid>/io - Display the IO accounting fields
+-------------------------------------------------------
+
+This file contains IO statistics for each running process
+
+Example
+-------
+
+test:/tmp # dd if=/dev/zero of=/tmp/test.dat &
+[1] 3828
+
+test:/tmp # cat /proc/3828/io
+rchar: 323934931
+wchar: 323929600
+syscr: 632687
+syscw: 632675
+read_bytes: 0
+write_bytes: 323932160
+cancelled_write_bytes: 0
+
+
+Description
+-----------
+
+rchar
+-----
+
+I/O counter: chars read
+The number of bytes which this task has caused to be read from storage. This
+is simply the sum of bytes which this process passed to read() and pread().
+It includes things like tty IO and it is unaffected by whether or not actual
+physical disk IO was required (the read might have been satisfied from
+pagecache)
+
+
+wchar
+-----
+
+I/O counter: chars written
+The number of bytes which this task has caused, or shall cause to be written
+to disk. Similar caveats apply here as with rchar.
+
+
+syscr
+-----
+
+I/O counter: read syscalls
+Attempt to count the number of read I/O operations, i.e. syscalls like read()
+and pread().
+
+
+syscw
+-----
+
+I/O counter: write syscalls
+Attempt to count the number of write I/O operations, i.e. syscalls like
+write() and pwrite().
+
+
+read_bytes
+----------
+
+I/O counter: bytes read
+Attempt to count the number of bytes which this process really did cause to
+be fetched from the storage layer. Done at the submit_bio() level, so it is
+accurate for block-backed filesystems. <please add status regarding NFS and
+CIFS at a later time>
+
+
+write_bytes
+-----------
+
+I/O counter: bytes written
+Attempt to count the number of bytes which this process caused to be sent to
+the storage layer. This is done at page-dirtying time.
+
+
+cancelled_write_bytes
+---------------------
+
+The big inaccuracy here is truncate. If a process writes 1MB to a file and
+then deletes the file, it will in fact perform no writeout. But it will have
+been accounted as having caused 1MB of write.
+In other words: The number of bytes which this process caused to not happen,
+by truncating pagecache. A task can cause "negative" IO too. If this task
+truncates some dirty pagecache, some IO which another task has been accounted
+for (in its write_bytes) will not be happening. We _could_ just subtract that
+from the truncating task's write_bytes, but there is information loss in doing
+that.
+
+
+Note
+----
+
+At its current implementation state, this is a bit racy on 32-bit machines: if
+process A reads process B's /proc/pid/io while process B is updating one of
+those 64-bit counters, process A could see an intermediate result.
+
+
+More information about this can be found within the taskstats documentation in
+Documentation/accounting.
+
+3.4 /proc/<pid>/coredump_filter - Core dump filtering settings
+---------------------------------------------------------------
+When a process is dumped, all anonymous memory is written to a core file as
+long as the size of the core file isn't limited. But sometimes we don't want
+to dump some memory segments, for example, huge shared memory or DAX.
+Conversely, sometimes we want to save file-backed memory segments into a core
+file, not only the individual files.
+
+/proc/<pid>/coredump_filter allows you to customize which memory segments
+will be dumped when the <pid> process is dumped. coredump_filter is a bitmask
+of memory types. If a bit of the bitmask is set, memory segments of the
+corresponding memory type are dumped, otherwise they are not dumped.
+
+The following 9 memory types are supported:
+  - (bit 0) anonymous private memory
+  - (bit 1) anonymous shared memory
+  - (bit 2) file-backed private memory
+  - (bit 3) file-backed shared memory
+  - (bit 4) ELF header pages in file-backed private memory areas (it is
+            effective only if the bit 2 is cleared)
+  - (bit 5) hugetlb private memory
+  - (bit 6) hugetlb shared memory
+  - (bit 7) DAX private memory
+  - (bit 8) DAX shared memory
+
+  Note that MMIO pages such as frame buffer are never dumped and vDSO pages
+  are always dumped regardless of the bitmask status.
+
+  Note that bits 0-4 don't affect hugetlb or DAX memory. hugetlb memory is
+  only affected by bit 5-6, and DAX is only affected by bits 7-8.
+
+The default value of coredump_filter is 0x33; this means all anonymous memory
+segments, ELF header pages and hugetlb private memory are dumped.
+
+If you don't want to dump all shared memory segments attached to pid 1234,
+write 0x31 to the process's proc file.
+
+  $ echo 0x31 > /proc/1234/coredump_filter
+
+When a new process is created, the process inherits the bitmask status from its
+parent. It is useful to set up coredump_filter before the program runs.
+For example:
+
+  $ echo 0x7 > /proc/self/coredump_filter
+  $ ./some_program
+
+3.5	/proc/<pid>/mountinfo - Information about mounts
+--------------------------------------------------------
+
+This file contains lines of the form:
+
+36 35 98:0 /mnt1 /mnt2 rw,noatime master:1 - ext3 /dev/root rw,errors=continue
+(1)(2)(3)   (4)   (5)      (6)      (7)   (8) (9)   (10)         (11)
+
+(1) mount ID:  unique identifier of the mount (may be reused after umount)
+(2) parent ID:  ID of parent (or of self for the top of the mount tree)
+(3) major:minor:  value of st_dev for files on filesystem
+(4) root:  root of the mount within the filesystem
+(5) mount point:  mount point relative to the process's root
+(6) mount options:  per mount options
+(7) optional fields:  zero or more fields of the form "tag[:value]"
+(8) separator:  marks the end of the optional fields
+(9) filesystem type:  name of filesystem of the form "type[.subtype]"
+(10) mount source:  filesystem specific information or "none"
+(11) super options:  per super block options
+
+Parsers should ignore all unrecognised optional fields.  Currently the
+possible optional fields are:
+
+shared:X  mount is shared in peer group X
+master:X  mount is slave to peer group X
+propagate_from:X  mount is slave and receives propagation from peer group X (*)
+unbindable  mount is unbindable
+
+(*) X is the closest dominant peer group under the process's root.  If
+X is the immediate master of the mount, or if there's no dominant peer
+group under the same root, then only the "master:X" field is present
+and not the "propagate_from:X" field.
+
+For more information on mount propagation see:
+
+  Documentation/filesystems/sharedsubtree.txt
+
+
+3.6	/proc/<pid>/comm  & /proc/<pid>/task/<tid>/comm
+--------------------------------------------------------
+These files provide a method to access a tasks comm value. It also allows for
+a task to set its own or one of its thread siblings comm value. The comm value
+is limited in size compared to the cmdline value, so writing anything longer
+then the kernel's TASK_COMM_LEN (currently 16 chars) will result in a truncated
+comm value.
+
+
+3.7	/proc/<pid>/task/<tid>/children - Information about task children
+-------------------------------------------------------------------------
+This file provides a fast way to retrieve first level children pids
+of a task pointed by <pid>/<tid> pair. The format is a space separated
+stream of pids.
+
+Note the "first level" here -- if a child has own children they will
+not be listed here, one needs to read /proc/<children-pid>/task/<tid>/children
+to obtain the descendants.
+
+Since this interface is intended to be fast and cheap it doesn't
+guarantee to provide precise results and some children might be
+skipped, especially if they've exited right after we printed their
+pids, so one need to either stop or freeze processes being inspected
+if precise results are needed.
+
+
+3.8	/proc/<pid>/fdinfo/<fd> - Information about opened file
+---------------------------------------------------------------
+This file provides information associated with an opened file. The regular
+files have at least three fields -- 'pos', 'flags' and mnt_id. The 'pos'
+represents the current offset of the opened file in decimal form [see lseek(2)
+for details], 'flags' denotes the octal O_xxx mask the file has been
+created with [see open(2) for details] and 'mnt_id' represents mount ID of
+the file system containing the opened file [see 3.5 /proc/<pid>/mountinfo
+for details].
+
+A typical output is
+
+	pos:	0
+	flags:	0100002
+	mnt_id:	19
+
+All locks associated with a file descriptor are shown in its fdinfo too.
+
+lock:       1: FLOCK  ADVISORY  WRITE 359 00:13:11691 0 EOF
+
+The files such as eventfd, fsnotify, signalfd, epoll among the regular pos/flags
+pair provide additional information particular to the objects they represent.
+
+	Eventfd files
+	~~~~~~~~~~~~~
+	pos:	0
+	flags:	04002
+	mnt_id:	9
+	eventfd-count:	5a
+
+	where 'eventfd-count' is hex value of a counter.
+
+	Signalfd files
+	~~~~~~~~~~~~~~
+	pos:	0
+	flags:	04002
+	mnt_id:	9
+	sigmask:	0000000000000200
+
+	where 'sigmask' is hex value of the signal mask associated
+	with a file.
+
+	Epoll files
+	~~~~~~~~~~~
+	pos:	0
+	flags:	02
+	mnt_id:	9
+	tfd:        5 events:       1d data: ffffffffffffffff pos:0 ino:61af sdev:7
+
+	where 'tfd' is a target file descriptor number in decimal form,
+	'events' is events mask being watched and the 'data' is data
+	associated with a target [see epoll(7) for more details].
+
+	The 'pos' is current offset of the target file in decimal form
+	[see lseek(2)], 'ino' and 'sdev' are inode and device numbers
+	where target file resides, all in hex format.
+
+	Fsnotify files
+	~~~~~~~~~~~~~~
+	For inotify files the format is the following
+
+	pos:	0
+	flags:	02000000
+	inotify wd:3 ino:9e7e sdev:800013 mask:800afce ignored_mask:0 fhandle-bytes:8 fhandle-type:1 f_handle:7e9e0000640d1b6d
+
+	where 'wd' is a watch descriptor in decimal form, ie a target file
+	descriptor number, 'ino' and 'sdev' are inode and device where the
+	target file resides and the 'mask' is the mask of events, all in hex
+	form [see inotify(7) for more details].
+
+	If the kernel was built with exportfs support, the path to the target
+	file is encoded as a file handle.  The file handle is provided by three
+	fields 'fhandle-bytes', 'fhandle-type' and 'f_handle', all in hex
+	format.
+
+	If the kernel is built without exportfs support the file handle won't be
+	printed out.
+
+	If there is no inotify mark attached yet the 'inotify' line will be omitted.
+
+	For fanotify files the format is
+
+	pos:	0
+	flags:	02
+	mnt_id:	9
+	fanotify flags:10 event-flags:0
+	fanotify mnt_id:12 mflags:40 mask:38 ignored_mask:40000003
+	fanotify ino:4f969 sdev:800013 mflags:0 mask:3b ignored_mask:40000000 fhandle-bytes:8 fhandle-type:1 f_handle:69f90400c275b5b4
+
+	where fanotify 'flags' and 'event-flags' are values used in fanotify_init
+	call, 'mnt_id' is the mount point identifier, 'mflags' is the value of
+	flags associated with mark which are tracked separately from events
+	mask. 'ino', 'sdev' are target inode and device, 'mask' is the events
+	mask and 'ignored_mask' is the mask of events which are to be ignored.
+	All in hex format. Incorporation of 'mflags', 'mask' and 'ignored_mask'
+	does provide information about flags and mask used in fanotify_mark
+	call [see fsnotify manpage for details].
+
+	While the first three lines are mandatory and always printed, the rest is
+	optional and may be omitted if no marks created yet.
+
+	Timerfd files
+	~~~~~~~~~~~~~
+
+	pos:	0
+	flags:	02
+	mnt_id:	9
+	clockid: 0
+	ticks: 0
+	settime flags: 01
+	it_value: (0, 49406829)
+	it_interval: (1, 0)
+
+	where 'clockid' is the clock type and 'ticks' is the number of the timer expirations
+	that have occurred [see timerfd_create(2) for details]. 'settime flags' are
+	flags in octal form been used to setup the timer [see timerfd_settime(2) for
+	details]. 'it_value' is remaining time until the timer exiration.
+	'it_interval' is the interval for the timer. Note the timer might be set up
+	with TIMER_ABSTIME option which will be shown in 'settime flags', but 'it_value'
+	still exhibits timer's remaining time.
+
+3.9	/proc/<pid>/map_files - Information about memory mapped files
+---------------------------------------------------------------------
+This directory contains symbolic links which represent memory mapped files
+the process is maintaining.  Example output:
+
+     | lr-------- 1 root root 64 Jan 27 11:24 333c600000-333c620000 -> /usr/lib64/ld-2.18.so
+     | lr-------- 1 root root 64 Jan 27 11:24 333c81f000-333c820000 -> /usr/lib64/ld-2.18.so
+     | lr-------- 1 root root 64 Jan 27 11:24 333c820000-333c821000 -> /usr/lib64/ld-2.18.so
+     | ...
+     | lr-------- 1 root root 64 Jan 27 11:24 35d0421000-35d0422000 -> /usr/lib64/libselinux.so.1
+     | lr-------- 1 root root 64 Jan 27 11:24 400000-41a000 -> /usr/bin/ls
+
+The name of a link represents the virtual memory bounds of a mapping, i.e.
+vm_area_struct::vm_start-vm_area_struct::vm_end.
+
+The main purpose of the map_files is to retrieve a set of memory mapped
+files in a fast way instead of parsing /proc/<pid>/maps or
+/proc/<pid>/smaps, both of which contain many more records.  At the same
+time one can open(2) mappings from the listings of two processes and
+comparing their inode numbers to figure out which anonymous memory areas
+are actually shared.
+
+3.10	/proc/<pid>/timerslack_ns - Task timerslack value
+---------------------------------------------------------
+This file provides the value of the task's timerslack value in nanoseconds.
+This value specifies a amount of time that normal timers may be deferred
+in order to coalesce timers and avoid unnecessary wakeups.
+
+This allows a task's interactivity vs power consumption trade off to be
+adjusted.
+
+Writing 0 to the file will set the tasks timerslack to the default value.
+
+Valid values are from 0 - ULLONG_MAX
+
+An application setting the value must have PTRACE_MODE_ATTACH_FSCREDS level
+permissions on the task specified to change its timerslack_ns value.
+
+3.11	/proc/<pid>/patch_state - Livepatch patch operation state
+-----------------------------------------------------------------
+When CONFIG_LIVEPATCH is enabled, this file displays the value of the
+patch state for the task.
+
+A value of '-1' indicates that no patch is in transition.
+
+A value of '0' indicates that a patch is in transition and the task is
+unpatched.  If the patch is being enabled, then the task hasn't been
+patched yet.  If the patch is being disabled, then the task has already
+been unpatched.
+
+A value of '1' indicates that a patch is in transition and the task is
+patched.  If the patch is being enabled, then the task has already been
+patched.  If the patch is being disabled, then the task hasn't been
+unpatched yet.
+
+
+------------------------------------------------------------------------------
+Configuring procfs
+------------------------------------------------------------------------------
+
+4.1	Mount options
+---------------------
+
+The following mount options are supported:
+
+	hidepid=	Set /proc/<pid>/ access mode.
+	gid=		Set the group authorized to learn processes information.
+
+hidepid=0 means classic mode - everybody may access all /proc/<pid>/ directories
+(default).
+
+hidepid=1 means users may not access any /proc/<pid>/ directories but their
+own.  Sensitive files like cmdline, sched*, status are now protected against
+other users.  This makes it impossible to learn whether any user runs
+specific program (given the program doesn't reveal itself by its behaviour).
+As an additional bonus, as /proc/<pid>/cmdline is unaccessible for other users,
+poorly written programs passing sensitive information via program arguments are
+now protected against local eavesdroppers.
+
+hidepid=2 means hidepid=1 plus all /proc/<pid>/ will be fully invisible to other
+users.  It doesn't mean that it hides a fact whether a process with a specific
+pid value exists (it can be learned by other means, e.g. by "kill -0 $PID"),
+but it hides process' uid and gid, which may be learned by stat()'ing
+/proc/<pid>/ otherwise.  It greatly complicates an intruder's task of gathering
+information about running processes, whether some daemon runs with elevated
+privileges, whether other user runs some sensitive program, whether other users
+run any program at all, etc.
+
+gid= defines a group authorized to learn processes information otherwise
+prohibited by hidepid=.  If you use some daemon like identd which needs to learn
+information about processes information, just add identd to this group.
diff --git a/Documentation/filesystems/qnx6.txt b/Documentation/filesystems/qnx6.txt
new file mode 100644
index 000000000..4f3d6a882
--- /dev/null
+++ b/Documentation/filesystems/qnx6.txt
@@ -0,0 +1,174 @@
+The QNX6 Filesystem
+===================
+
+The qnx6fs is used by newer QNX operating system versions. (e.g. Neutrino)
+It got introduced in QNX 6.4.0 and is used default since 6.4.1.
+
+Option
+======
+
+mmi_fs		Mount filesystem as used for example by Audi MMI 3G system
+
+Specification
+=============
+
+qnx6fs shares many properties with traditional Unix filesystems. It has the
+concepts of blocks, inodes and directories.
+On QNX it is possible to create little endian and big endian qnx6 filesystems.
+This feature makes it possible to create and use a different endianness fs
+for the target (QNX is used on quite a range of embedded systems) platform
+running on a different endianness.
+The Linux driver handles endianness transparently. (LE and BE)
+
+Blocks
+------
+
+The space in the device or file is split up into blocks. These are a fixed
+size of 512, 1024, 2048 or 4096, which is decided when the filesystem is
+created.
+Blockpointers are 32bit, so the maximum space that can be addressed is
+2^32 * 4096 bytes or 16TB
+
+The superblocks
+---------------
+
+The superblock contains all global information about the filesystem.
+Each qnx6fs got two superblocks, each one having a 64bit serial number.
+That serial number is used to identify the "active" superblock.
+In write mode with reach new snapshot (after each synchronous write), the
+serial of the new master superblock is increased (old superblock serial + 1)
+
+So basically the snapshot functionality is realized by an atomic final
+update of the serial number. Before updating that serial, all modifications
+are done by copying all modified blocks during that specific write request
+(or period) and building up a new (stable) filesystem structure under the
+inactive superblock.
+
+Each superblock holds a set of root inodes for the different filesystem
+parts. (Inode, Bitmap and Longfilenames)
+Each of these root nodes holds information like total size of the stored
+data and the addressing levels in that specific tree.
+If the level value is 0, up to 16 direct blocks can be addressed by each
+node.
+Level 1 adds an additional indirect addressing level where each indirect
+addressing block holds up to blocksize / 4 bytes pointers to data blocks.
+Level 2 adds an additional indirect addressing block level (so, already up
+to 16 * 256 * 256 = 1048576 blocks that can be addressed by such a tree).
+
+Unused block pointers are always set to ~0 - regardless of root node,
+indirect addressing blocks or inodes.
+Data leaves are always on the lowest level. So no data is stored on upper
+tree levels.
+
+The first Superblock is located at 0x2000. (0x2000 is the bootblock size)
+The Audi MMI 3G first superblock directly starts at byte 0.
+Second superblock position can either be calculated from the superblock
+information (total number of filesystem blocks) or by taking the highest
+device address, zeroing the last 3 bytes and then subtracting 0x1000 from
+that address.
+
+0x1000 is the size reserved for each superblock - regardless of the
+blocksize of the filesystem.
+
+Inodes
+------
+
+Each object in the filesystem is represented by an inode. (index node)
+The inode structure contains pointers to the filesystem blocks which contain
+the data held in the object and all of the metadata about an object except
+its longname. (filenames longer than 27 characters)
+The metadata about an object includes the permissions, owner, group, flags,
+size, number of blocks used, access time, change time and modification time.
+
+Object mode field is POSIX format. (which makes things easier)
+
+There are also pointers to the first 16 blocks, if the object data can be
+addressed with 16 direct blocks.
+For more than 16 blocks an indirect addressing in form of another tree is
+used. (scheme is the same as the one used for the superblock root nodes)
+
+The filesize is stored 64bit. Inode counting starts with 1. (whilst long
+filename inodes start with 0)
+
+Directories
+-----------
+
+A directory is a filesystem object and has an inode just like a file.
+It is a specially formatted file containing records which associate each
+name with an inode number.
+'.' inode number points to the directory inode
+'..' inode number points to the parent directory inode
+Eeach filename record additionally got a filename length field.
+
+One special case are long filenames or subdirectory names.
+These got set a filename length field of 0xff in the corresponding directory
+record plus the longfile inode number also stored in that record.
+With that longfilename inode number, the longfilename tree can be walked
+starting with the superblock longfilename root node pointers.
+
+Special files
+-------------
+
+Symbolic links are also filesystem objects with inodes. They got a specific
+bit in the inode mode field identifying them as symbolic link.
+The directory entry file inode pointer points to the target file inode.
+
+Hard links got an inode, a directory entry, but a specific mode bit set,
+no block pointers and the directory file record pointing to the target file
+inode.
+
+Character and block special devices do not exist in QNX as those files
+are handled by the QNX kernel/drivers and created in /dev independent of the
+underlaying filesystem.
+
+Long filenames
+--------------
+
+Long filenames are stored in a separate addressing tree. The staring point
+is the longfilename root node in the active superblock.
+Each data block (tree leaves) holds one long filename. That filename is
+limited to 510 bytes. The first two starting bytes are used as length field
+for the actual filename.
+If that structure shall fit for all allowed blocksizes, it is clear why there
+is a limit of 510 bytes for the actual filename stored.
+
+Bitmap
+------
+
+The qnx6fs filesystem allocation bitmap is stored in a tree under bitmap
+root node in the superblock and each bit in the bitmap represents one
+filesystem block.
+The first block is block 0, which starts 0x1000 after superblock start.
+So for a normal qnx6fs 0x3000 (bootblock + superblock) is the physical
+address at which block 0 is located.
+
+Bits at the end of the last bitmap block are set to 1, if the device is
+smaller than addressing space in the bitmap.
+
+Bitmap system area
+------------------
+
+The bitmap itself is divided into three parts.
+First the system area, that is split into two halves.
+Then userspace.
+
+The requirement for a static, fixed preallocated system area comes from how
+qnx6fs deals with writes.
+Each superblock got it's own half of the system area. So superblock #1
+always uses blocks from the lower half whilst superblock #2 just writes to
+blocks represented by the upper half bitmap system area bits.
+
+Bitmap blocks, Inode blocks and indirect addressing blocks for those two
+tree structures are treated as system blocks.
+
+The rational behind that is that a write request can work on a new snapshot
+(system area of the inactive - resp. lower serial numbered superblock) while
+at the same time there is still a complete stable filesystem structer in the
+other half of the system area.
+
+When finished with writing (a sync write is completed, the maximum sync leap
+time or a filesystem sync is requested), serial of the previously inactive
+superblock atomically is increased and the fs switches over to that - then
+stable declared - superblock.
+
+For all data outside the system area, blocks are just copied while writing.
diff --git a/Documentation/filesystems/quota.txt b/Documentation/filesystems/quota.txt
new file mode 100644
index 000000000..32874b06e
--- /dev/null
+++ b/Documentation/filesystems/quota.txt
@@ -0,0 +1,68 @@
+
+Quota subsystem
+===============
+
+Quota subsystem allows system administrator to set limits on used space and
+number of used inodes (inode is a filesystem structure which is associated with
+each file or directory) for users and/or groups. For both used space and number
+of used inodes there are actually two limits. The first one is called softlimit
+and the second one hardlimit.  A user can never exceed a hardlimit for any
+resource (unless he has CAP_SYS_RESOURCE capability). User is allowed to exceed
+softlimit but only for limited period of time. This period is called "grace
+period" or "grace time". When grace time is over, user is not able to allocate
+more space/inodes until he frees enough of them to get below softlimit.
+
+Quota limits (and amount of grace time) are set independently for each
+filesystem.
+
+For more details about quota design, see the documentation in quota-tools package
+(http://sourceforge.net/projects/linuxquota).
+
+Quota netlink interface
+=======================
+When user exceeds a softlimit, runs out of grace time or reaches hardlimit,
+quota subsystem traditionally printed a message to the controlling terminal of
+the process which caused the excess. This method has the disadvantage that
+when user is using a graphical desktop he usually cannot see the message.
+Thus quota netlink interface has been designed to pass information about
+the above events to userspace. There they can be captured by an application
+and processed accordingly.
+
+The interface uses generic netlink framework (see
+http://lwn.net/Articles/208755/ and http://people.suug.ch/~tgr/libnl/ for more
+details about this layer). The name of the quota generic netlink interface
+is "VFS_DQUOT". Definitions of constants below are in <linux/quota.h>.
+Since the quota netlink protocol is not namespace aware, quota netlink messages
+are sent only in initial network namespace.
+
+Currently, the interface supports only one message type QUOTA_NL_C_WARNING.
+This command is used to send a notification about any of the above mentioned
+events. Each message has six attributes. These are (type of the argument is
+in parentheses):
+        QUOTA_NL_A_QTYPE (u32)
+	  - type of quota being exceeded (one of USRQUOTA, GRPQUOTA)
+        QUOTA_NL_A_EXCESS_ID (u64)
+	  - UID/GID (depends on quota type) of user / group whose limit
+	    is being exceeded.
+        QUOTA_NL_A_CAUSED_ID (u64)
+	  - UID of a user who caused the event
+        QUOTA_NL_A_WARNING (u32)
+	  - what kind of limit is exceeded:
+		QUOTA_NL_IHARDWARN - inode hardlimit
+		QUOTA_NL_ISOFTLONGWARN - inode softlimit is exceeded longer
+		  than given grace period
+		QUOTA_NL_ISOFTWARN - inode softlimit
+		QUOTA_NL_BHARDWARN - space (block) hardlimit
+		QUOTA_NL_BSOFTLONGWARN - space (block) softlimit is exceeded
+		  longer than given grace period.
+		QUOTA_NL_BSOFTWARN - space (block) softlimit
+	  - four warnings are also defined for the event when user stops
+	    exceeding some limit:
+		QUOTA_NL_IHARDBELOW - inode hardlimit
+		QUOTA_NL_ISOFTBELOW - inode softlimit
+		QUOTA_NL_BHARDBELOW - space (block) hardlimit
+		QUOTA_NL_BSOFTBELOW - space (block) softlimit
+        QUOTA_NL_A_DEV_MAJOR (u32)
+	  - major number of a device with the affected filesystem
+        QUOTA_NL_A_DEV_MINOR (u32)
+	  - minor number of a device with the affected filesystem
diff --git a/Documentation/filesystems/ramfs-rootfs-initramfs.txt b/Documentation/filesystems/ramfs-rootfs-initramfs.txt
new file mode 100644
index 000000000..79637d227
--- /dev/null
+++ b/Documentation/filesystems/ramfs-rootfs-initramfs.txt
@@ -0,0 +1,359 @@
+ramfs, rootfs and initramfs
+October 17, 2005
+Rob Landley <rob@landley.net>
+=============================
+
+What is ramfs?
+--------------
+
+Ramfs is a very simple filesystem that exports Linux's disk caching
+mechanisms (the page cache and dentry cache) as a dynamically resizable
+RAM-based filesystem.
+
+Normally all files are cached in memory by Linux.  Pages of data read from
+backing store (usually the block device the filesystem is mounted on) are kept
+around in case it's needed again, but marked as clean (freeable) in case the
+Virtual Memory system needs the memory for something else.  Similarly, data
+written to files is marked clean as soon as it has been written to backing
+store, but kept around for caching purposes until the VM reallocates the
+memory.  A similar mechanism (the dentry cache) greatly speeds up access to
+directories.
+
+With ramfs, there is no backing store.  Files written into ramfs allocate
+dentries and page cache as usual, but there's nowhere to write them to.
+This means the pages are never marked clean, so they can't be freed by the
+VM when it's looking to recycle memory.
+
+The amount of code required to implement ramfs is tiny, because all the
+work is done by the existing Linux caching infrastructure.  Basically,
+you're mounting the disk cache as a filesystem.  Because of this, ramfs is not
+an optional component removable via menuconfig, since there would be negligible
+space savings.
+
+ramfs and ramdisk:
+------------------
+
+The older "ram disk" mechanism created a synthetic block device out of
+an area of RAM and used it as backing store for a filesystem.  This block
+device was of fixed size, so the filesystem mounted on it was of fixed
+size.  Using a ram disk also required unnecessarily copying memory from the
+fake block device into the page cache (and copying changes back out), as well
+as creating and destroying dentries.  Plus it needed a filesystem driver
+(such as ext2) to format and interpret this data.
+
+Compared to ramfs, this wastes memory (and memory bus bandwidth), creates
+unnecessary work for the CPU, and pollutes the CPU caches.  (There are tricks
+to avoid this copying by playing with the page tables, but they're unpleasantly
+complicated and turn out to be about as expensive as the copying anyway.)
+More to the point, all the work ramfs is doing has to happen _anyway_,
+since all file access goes through the page and dentry caches.  The RAM
+disk is simply unnecessary; ramfs is internally much simpler.
+
+Another reason ramdisks are semi-obsolete is that the introduction of
+loopback devices offered a more flexible and convenient way to create
+synthetic block devices, now from files instead of from chunks of memory.
+See losetup (8) for details.
+
+ramfs and tmpfs:
+----------------
+
+One downside of ramfs is you can keep writing data into it until you fill
+up all memory, and the VM can't free it because the VM thinks that files
+should get written to backing store (rather than swap space), but ramfs hasn't
+got any backing store.  Because of this, only root (or a trusted user) should
+be allowed write access to a ramfs mount.
+
+A ramfs derivative called tmpfs was created to add size limits, and the ability
+to write the data to swap space.  Normal users can be allowed write access to
+tmpfs mounts.  See Documentation/filesystems/tmpfs.txt for more information.
+
+What is rootfs?
+---------------
+
+Rootfs is a special instance of ramfs (or tmpfs, if that's enabled), which is
+always present in 2.6 systems.  You can't unmount rootfs for approximately the
+same reason you can't kill the init process; rather than having special code
+to check for and handle an empty list, it's smaller and simpler for the kernel
+to just make sure certain lists can't become empty.
+
+Most systems just mount another filesystem over rootfs and ignore it.  The
+amount of space an empty instance of ramfs takes up is tiny.
+
+If CONFIG_TMPFS is enabled, rootfs will use tmpfs instead of ramfs by
+default.  To force ramfs, add "rootfstype=ramfs" to the kernel command
+line.
+
+What is initramfs?
+------------------
+
+All 2.6 Linux kernels contain a gzipped "cpio" format archive, which is
+extracted into rootfs when the kernel boots up.  After extracting, the kernel
+checks to see if rootfs contains a file "init", and if so it executes it as PID
+1.  If found, this init process is responsible for bringing the system the
+rest of the way up, including locating and mounting the real root device (if
+any).  If rootfs does not contain an init program after the embedded cpio
+archive is extracted into it, the kernel will fall through to the older code
+to locate and mount a root partition, then exec some variant of /sbin/init
+out of that.
+
+All this differs from the old initrd in several ways:
+
+  - The old initrd was always a separate file, while the initramfs archive is
+    linked into the linux kernel image.  (The directory linux-*/usr is devoted
+    to generating this archive during the build.)
+
+  - The old initrd file was a gzipped filesystem image (in some file format,
+    such as ext2, that needed a driver built into the kernel), while the new
+    initramfs archive is a gzipped cpio archive (like tar only simpler,
+    see cpio(1) and Documentation/early-userspace/buffer-format.txt).  The
+    kernel's cpio extraction code is not only extremely small, it's also
+    __init text and data that can be discarded during the boot process.
+
+  - The program run by the old initrd (which was called /initrd, not /init) did
+    some setup and then returned to the kernel, while the init program from
+    initramfs is not expected to return to the kernel.  (If /init needs to hand
+    off control it can overmount / with a new root device and exec another init
+    program.  See the switch_root utility, below.)
+
+  - When switching another root device, initrd would pivot_root and then
+    umount the ramdisk.  But initramfs is rootfs: you can neither pivot_root
+    rootfs, nor unmount it.  Instead delete everything out of rootfs to
+    free up the space (find -xdev / -exec rm '{}' ';'), overmount rootfs
+    with the new root (cd /newmount; mount --move . /; chroot .), attach
+    stdin/stdout/stderr to the new /dev/console, and exec the new init.
+
+    Since this is a remarkably persnickety process (and involves deleting
+    commands before you can run them), the klibc package introduced a helper
+    program (utils/run_init.c) to do all this for you.  Most other packages
+    (such as busybox) have named this command "switch_root".
+
+Populating initramfs:
+---------------------
+
+The 2.6 kernel build process always creates a gzipped cpio format initramfs
+archive and links it into the resulting kernel binary.  By default, this
+archive is empty (consuming 134 bytes on x86).
+
+The config option CONFIG_INITRAMFS_SOURCE (in General Setup in menuconfig,
+and living in usr/Kconfig) can be used to specify a source for the
+initramfs archive, which will automatically be incorporated into the
+resulting binary.  This option can point to an existing gzipped cpio
+archive, a directory containing files to be archived, or a text file
+specification such as the following example:
+
+  dir /dev 755 0 0
+  nod /dev/console 644 0 0 c 5 1
+  nod /dev/loop0 644 0 0 b 7 0
+  dir /bin 755 1000 1000
+  slink /bin/sh busybox 777 0 0
+  file /bin/busybox initramfs/busybox 755 0 0
+  dir /proc 755 0 0
+  dir /sys 755 0 0
+  dir /mnt 755 0 0
+  file /init initramfs/init.sh 755 0 0
+
+Run "usr/gen_init_cpio" (after the kernel build) to get a usage message
+documenting the above file format.
+
+One advantage of the configuration file is that root access is not required to
+set permissions or create device nodes in the new archive.  (Note that those
+two example "file" entries expect to find files named "init.sh" and "busybox" in
+a directory called "initramfs", under the linux-2.6.* directory.  See
+Documentation/early-userspace/README for more details.)
+
+The kernel does not depend on external cpio tools.  If you specify a
+directory instead of a configuration file, the kernel's build infrastructure
+creates a configuration file from that directory (usr/Makefile calls
+usr/gen_initramfs_list.sh), and proceeds to package up that directory
+using the config file (by feeding it to usr/gen_init_cpio, which is created
+from usr/gen_init_cpio.c).  The kernel's build-time cpio creation code is
+entirely self-contained, and the kernel's boot-time extractor is also
+(obviously) self-contained.
+
+The one thing you might need external cpio utilities installed for is creating
+or extracting your own preprepared cpio files to feed to the kernel build
+(instead of a config file or directory).
+
+The following command line can extract a cpio image (either by the above script
+or by the kernel build) back into its component files:
+
+  cpio -i -d -H newc -F initramfs_data.cpio --no-absolute-filenames
+
+The following shell script can create a prebuilt cpio archive you can
+use in place of the above config file:
+
+  #!/bin/sh
+
+  # Copyright 2006 Rob Landley <rob@landley.net> and TimeSys Corporation.
+  # Licensed under GPL version 2
+
+  if [ $# -ne 2 ]
+  then
+    echo "usage: mkinitramfs directory imagename.cpio.gz"
+    exit 1
+  fi
+
+  if [ -d "$1" ]
+  then
+    echo "creating $2 from $1"
+    (cd "$1"; find . | cpio -o -H newc | gzip) > "$2"
+  else
+    echo "First argument must be a directory"
+    exit 1
+  fi
+
+Note: The cpio man page contains some bad advice that will break your initramfs
+archive if you follow it.  It says "A typical way to generate the list
+of filenames is with the find command; you should give find the -depth option
+to minimize problems with permissions on directories that are unwritable or not
+searchable."  Don't do this when creating initramfs.cpio.gz images, it won't
+work.  The Linux kernel cpio extractor won't create files in a directory that
+doesn't exist, so the directory entries must go before the files that go in
+those directories.  The above script gets them in the right order.
+
+External initramfs images:
+--------------------------
+
+If the kernel has initrd support enabled, an external cpio.gz archive can also
+be passed into a 2.6 kernel in place of an initrd.  In this case, the kernel
+will autodetect the type (initramfs, not initrd) and extract the external cpio
+archive into rootfs before trying to run /init.
+
+This has the memory efficiency advantages of initramfs (no ramdisk block
+device) but the separate packaging of initrd (which is nice if you have
+non-GPL code you'd like to run from initramfs, without conflating it with
+the GPL licensed Linux kernel binary).
+
+It can also be used to supplement the kernel's built-in initramfs image.  The
+files in the external archive will overwrite any conflicting files in
+the built-in initramfs archive.  Some distributors also prefer to customize
+a single kernel image with task-specific initramfs images, without recompiling.
+
+Contents of initramfs:
+----------------------
+
+An initramfs archive is a complete self-contained root filesystem for Linux.
+If you don't already understand what shared libraries, devices, and paths
+you need to get a minimal root filesystem up and running, here are some
+references:
+http://www.tldp.org/HOWTO/Bootdisk-HOWTO/
+http://www.tldp.org/HOWTO/From-PowerUp-To-Bash-Prompt-HOWTO.html
+http://www.linuxfromscratch.org/lfs/view/stable/
+
+The "klibc" package (http://www.kernel.org/pub/linux/libs/klibc) is
+designed to be a tiny C library to statically link early userspace
+code against, along with some related utilities.  It is BSD licensed.
+
+I use uClibc (http://www.uclibc.org) and busybox (http://www.busybox.net)
+myself.  These are LGPL and GPL, respectively.  (A self-contained initramfs
+package is planned for the busybox 1.3 release.)
+
+In theory you could use glibc, but that's not well suited for small embedded
+uses like this.  (A "hello world" program statically linked against glibc is
+over 400k.  With uClibc it's 7k.  Also note that glibc dlopens libnss to do
+name lookups, even when otherwise statically linked.)
+
+A good first step is to get initramfs to run a statically linked "hello world"
+program as init, and test it under an emulator like qemu (www.qemu.org) or
+User Mode Linux, like so:
+
+  cat > hello.c << EOF
+  #include <stdio.h>
+  #include <unistd.h>
+
+  int main(int argc, char *argv[])
+  {
+    printf("Hello world!\n");
+    sleep(999999999);
+  }
+  EOF
+  gcc -static hello.c -o init
+  echo init | cpio -o -H newc | gzip > test.cpio.gz
+  # Testing external initramfs using the initrd loading mechanism.
+  qemu -kernel /boot/vmlinuz -initrd test.cpio.gz /dev/zero
+
+When debugging a normal root filesystem, it's nice to be able to boot with
+"init=/bin/sh".  The initramfs equivalent is "rdinit=/bin/sh", and it's
+just as useful.
+
+Why cpio rather than tar?
+-------------------------
+
+This decision was made back in December, 2001.  The discussion started here:
+
+  http://www.uwsg.iu.edu/hypermail/linux/kernel/0112.2/1538.html
+
+And spawned a second thread (specifically on tar vs cpio), starting here:
+
+  http://www.uwsg.iu.edu/hypermail/linux/kernel/0112.2/1587.html
+
+The quick and dirty summary version (which is no substitute for reading
+the above threads) is:
+
+1) cpio is a standard.  It's decades old (from the AT&T days), and already
+   widely used on Linux (inside RPM, Red Hat's device driver disks).  Here's
+   a Linux Journal article about it from 1996:
+
+      http://www.linuxjournal.com/article/1213
+
+   It's not as popular as tar because the traditional cpio command line tools
+   require _truly_hideous_ command line arguments.  But that says nothing
+   either way about the archive format, and there are alternative tools,
+   such as:
+
+     http://freecode.com/projects/afio
+
+2) The cpio archive format chosen by the kernel is simpler and cleaner (and
+   thus easier to create and parse) than any of the (literally dozens of)
+   various tar archive formats.  The complete initramfs archive format is
+   explained in buffer-format.txt, created in usr/gen_init_cpio.c, and
+   extracted in init/initramfs.c.  All three together come to less than 26k
+   total of human-readable text.
+
+3) The GNU project standardizing on tar is approximately as relevant as
+   Windows standardizing on zip.  Linux is not part of either, and is free
+   to make its own technical decisions.
+
+4) Since this is a kernel internal format, it could easily have been
+   something brand new.  The kernel provides its own tools to create and
+   extract this format anyway.  Using an existing standard was preferable,
+   but not essential.
+
+5) Al Viro made the decision (quote: "tar is ugly as hell and not going to be
+   supported on the kernel side"):
+
+      http://www.uwsg.iu.edu/hypermail/linux/kernel/0112.2/1540.html
+
+   explained his reasoning:
+
+      http://www.uwsg.iu.edu/hypermail/linux/kernel/0112.2/1550.html
+      http://www.uwsg.iu.edu/hypermail/linux/kernel/0112.2/1638.html
+
+   and, most importantly, designed and implemented the initramfs code.
+
+Future directions:
+------------------
+
+Today (2.6.16), initramfs is always compiled in, but not always used.  The
+kernel falls back to legacy boot code that is reached only if initramfs does
+not contain an /init program.  The fallback is legacy code, there to ensure a
+smooth transition and allowing early boot functionality to gradually move to
+"early userspace" (I.E. initramfs).
+
+The move to early userspace is necessary because finding and mounting the real
+root device is complex.  Root partitions can span multiple devices (raid or
+separate journal).  They can be out on the network (requiring dhcp, setting a
+specific MAC address, logging into a server, etc).  They can live on removable
+media, with dynamically allocated major/minor numbers and persistent naming
+issues requiring a full udev implementation to sort out.  They can be
+compressed, encrypted, copy-on-write, loopback mounted, strangely partitioned,
+and so on.
+
+This kind of complexity (which inevitably includes policy) is rightly handled
+in userspace.  Both klibc and busybox/uClibc are working on simple initramfs
+packages to drop into a kernel build.
+
+The klibc package has now been accepted into Andrew Morton's 2.6.17-mm tree.
+The kernel's current early boot code (partition detection, etc) will probably
+be migrated into a default initramfs, automatically created and used by the
+kernel build.
diff --git a/Documentation/filesystems/relay.txt b/Documentation/filesystems/relay.txt
new file mode 100644
index 000000000..cd709a94d
--- /dev/null
+++ b/Documentation/filesystems/relay.txt
@@ -0,0 +1,494 @@
+relay interface (formerly relayfs)
+==================================
+
+The relay interface provides a means for kernel applications to
+efficiently log and transfer large quantities of data from the kernel
+to userspace via user-defined 'relay channels'.
+
+A 'relay channel' is a kernel->user data relay mechanism implemented
+as a set of per-cpu kernel buffers ('channel buffers'), each
+represented as a regular file ('relay file') in user space.  Kernel
+clients write into the channel buffers using efficient write
+functions; these automatically log into the current cpu's channel
+buffer.  User space applications mmap() or read() from the relay files
+and retrieve the data as it becomes available.  The relay files
+themselves are files created in a host filesystem, e.g. debugfs, and
+are associated with the channel buffers using the API described below.
+
+The format of the data logged into the channel buffers is completely
+up to the kernel client; the relay interface does however provide
+hooks which allow kernel clients to impose some structure on the
+buffer data.  The relay interface doesn't implement any form of data
+filtering - this also is left to the kernel client.  The purpose is to
+keep things as simple as possible.
+
+This document provides an overview of the relay interface API.  The
+details of the function parameters are documented along with the
+functions in the relay interface code - please see that for details.
+
+Semantics
+=========
+
+Each relay channel has one buffer per CPU, each buffer has one or more
+sub-buffers.  Messages are written to the first sub-buffer until it is
+too full to contain a new message, in which case it is written to
+the next (if available).  Messages are never split across sub-buffers.
+At this point, userspace can be notified so it empties the first
+sub-buffer, while the kernel continues writing to the next.
+
+When notified that a sub-buffer is full, the kernel knows how many
+bytes of it are padding i.e. unused space occurring because a complete
+message couldn't fit into a sub-buffer.  Userspace can use this
+knowledge to copy only valid data.
+
+After copying it, userspace can notify the kernel that a sub-buffer
+has been consumed.
+
+A relay channel can operate in a mode where it will overwrite data not
+yet collected by userspace, and not wait for it to be consumed.
+
+The relay channel itself does not provide for communication of such
+data between userspace and kernel, allowing the kernel side to remain
+simple and not impose a single interface on userspace.  It does
+provide a set of examples and a separate helper though, described
+below.
+
+The read() interface both removes padding and internally consumes the
+read sub-buffers; thus in cases where read(2) is being used to drain
+the channel buffers, special-purpose communication between kernel and
+user isn't necessary for basic operation.
+
+One of the major goals of the relay interface is to provide a low
+overhead mechanism for conveying kernel data to userspace.  While the
+read() interface is easy to use, it's not as efficient as the mmap()
+approach; the example code attempts to make the tradeoff between the
+two approaches as small as possible.
+
+klog and relay-apps example code
+================================
+
+The relay interface itself is ready to use, but to make things easier,
+a couple simple utility functions and a set of examples are provided.
+
+The relay-apps example tarball, available on the relay sourceforge
+site, contains a set of self-contained examples, each consisting of a
+pair of .c files containing boilerplate code for each of the user and
+kernel sides of a relay application.  When combined these two sets of
+boilerplate code provide glue to easily stream data to disk, without
+having to bother with mundane housekeeping chores.
+
+The 'klog debugging functions' patch (klog.patch in the relay-apps
+tarball) provides a couple of high-level logging functions to the
+kernel which allow writing formatted text or raw data to a channel,
+regardless of whether a channel to write into exists or not, or even
+whether the relay interface is compiled into the kernel or not.  These
+functions allow you to put unconditional 'trace' statements anywhere
+in the kernel or kernel modules; only when there is a 'klog handler'
+registered will data actually be logged (see the klog and kleak
+examples for details).
+
+It is of course possible to use the relay interface from scratch,
+i.e. without using any of the relay-apps example code or klog, but
+you'll have to implement communication between userspace and kernel,
+allowing both to convey the state of buffers (full, empty, amount of
+padding).  The read() interface both removes padding and internally
+consumes the read sub-buffers; thus in cases where read(2) is being
+used to drain the channel buffers, special-purpose communication
+between kernel and user isn't necessary for basic operation.  Things
+such as buffer-full conditions would still need to be communicated via
+some channel though.
+
+klog and the relay-apps examples can be found in the relay-apps
+tarball on http://relayfs.sourceforge.net
+
+The relay interface user space API
+==================================
+
+The relay interface implements basic file operations for user space
+access to relay channel buffer data.  Here are the file operations
+that are available and some comments regarding their behavior:
+
+open()	    enables user to open an _existing_ channel buffer.
+
+mmap()      results in channel buffer being mapped into the caller's
+	    memory space. Note that you can't do a partial mmap - you
+	    must map the entire file, which is NRBUF * SUBBUFSIZE.
+
+read()      read the contents of a channel buffer.  The bytes read are
+	    'consumed' by the reader, i.e. they won't be available
+	    again to subsequent reads.  If the channel is being used
+	    in no-overwrite mode (the default), it can be read at any
+	    time even if there's an active kernel writer.  If the
+	    channel is being used in overwrite mode and there are
+	    active channel writers, results may be unpredictable -
+	    users should make sure that all logging to the channel has
+	    ended before using read() with overwrite mode.  Sub-buffer
+	    padding is automatically removed and will not be seen by
+	    the reader.
+
+sendfile()  transfer data from a channel buffer to an output file
+	    descriptor. Sub-buffer padding is automatically removed
+	    and will not be seen by the reader.
+
+poll()      POLLIN/POLLRDNORM/POLLERR supported.  User applications are
+	    notified when sub-buffer boundaries are crossed.
+
+close()     decrements the channel buffer's refcount.  When the refcount
+	    reaches 0, i.e. when no process or kernel client has the
+	    buffer open, the channel buffer is freed.
+
+In order for a user application to make use of relay files, the
+host filesystem must be mounted.  For example,
+
+	mount -t debugfs debugfs /sys/kernel/debug
+
+NOTE:   the host filesystem doesn't need to be mounted for kernel
+	clients to create or use channels - it only needs to be
+	mounted when user space applications need access to the buffer
+	data.
+
+
+The relay interface kernel API
+==============================
+
+Here's a summary of the API the relay interface provides to in-kernel clients:
+
+TBD(curr. line MT:/API/)
+  channel management functions:
+
+    relay_open(base_filename, parent, subbuf_size, n_subbufs,
+               callbacks, private_data)
+    relay_close(chan)
+    relay_flush(chan)
+    relay_reset(chan)
+
+  channel management typically called on instigation of userspace:
+
+    relay_subbufs_consumed(chan, cpu, subbufs_consumed)
+
+  write functions:
+
+    relay_write(chan, data, length)
+    __relay_write(chan, data, length)
+    relay_reserve(chan, length)
+
+  callbacks:
+
+    subbuf_start(buf, subbuf, prev_subbuf, prev_padding)
+    buf_mapped(buf, filp)
+    buf_unmapped(buf, filp)
+    create_buf_file(filename, parent, mode, buf, is_global)
+    remove_buf_file(dentry)
+
+  helper functions:
+
+    relay_buf_full(buf)
+    subbuf_start_reserve(buf, length)
+
+
+Creating a channel
+------------------
+
+relay_open() is used to create a channel, along with its per-cpu
+channel buffers.  Each channel buffer will have an associated file
+created for it in the host filesystem, which can be and mmapped or
+read from in user space.  The files are named basename0...basenameN-1
+where N is the number of online cpus, and by default will be created
+in the root of the filesystem (if the parent param is NULL).  If you
+want a directory structure to contain your relay files, you should
+create it using the host filesystem's directory creation function,
+e.g. debugfs_create_dir(), and pass the parent directory to
+relay_open().  Users are responsible for cleaning up any directory
+structure they create, when the channel is closed - again the host
+filesystem's directory removal functions should be used for that,
+e.g. debugfs_remove().
+
+In order for a channel to be created and the host filesystem's files
+associated with its channel buffers, the user must provide definitions
+for two callback functions, create_buf_file() and remove_buf_file().
+create_buf_file() is called once for each per-cpu buffer from
+relay_open() and allows the user to create the file which will be used
+to represent the corresponding channel buffer.  The callback should
+return the dentry of the file created to represent the channel buffer.
+remove_buf_file() must also be defined; it's responsible for deleting
+the file(s) created in create_buf_file() and is called during
+relay_close().
+
+Here are some typical definitions for these callbacks, in this case
+using debugfs:
+
+/*
+ * create_buf_file() callback.  Creates relay file in debugfs.
+ */
+static struct dentry *create_buf_file_handler(const char *filename,
+                                              struct dentry *parent,
+                                              umode_t mode,
+                                              struct rchan_buf *buf,
+                                              int *is_global)
+{
+        return debugfs_create_file(filename, mode, parent, buf,
+	                           &relay_file_operations);
+}
+
+/*
+ * remove_buf_file() callback.  Removes relay file from debugfs.
+ */
+static int remove_buf_file_handler(struct dentry *dentry)
+{
+        debugfs_remove(dentry);
+
+        return 0;
+}
+
+/*
+ * relay interface callbacks
+ */
+static struct rchan_callbacks relay_callbacks =
+{
+        .create_buf_file = create_buf_file_handler,
+        .remove_buf_file = remove_buf_file_handler,
+};
+
+And an example relay_open() invocation using them:
+
+  chan = relay_open("cpu", NULL, SUBBUF_SIZE, N_SUBBUFS, &relay_callbacks, NULL);
+
+If the create_buf_file() callback fails, or isn't defined, channel
+creation and thus relay_open() will fail.
+
+The total size of each per-cpu buffer is calculated by multiplying the
+number of sub-buffers by the sub-buffer size passed into relay_open().
+The idea behind sub-buffers is that they're basically an extension of
+double-buffering to N buffers, and they also allow applications to
+easily implement random-access-on-buffer-boundary schemes, which can
+be important for some high-volume applications.  The number and size
+of sub-buffers is completely dependent on the application and even for
+the same application, different conditions will warrant different
+values for these parameters at different times.  Typically, the right
+values to use are best decided after some experimentation; in general,
+though, it's safe to assume that having only 1 sub-buffer is a bad
+idea - you're guaranteed to either overwrite data or lose events
+depending on the channel mode being used.
+
+The create_buf_file() implementation can also be defined in such a way
+as to allow the creation of a single 'global' buffer instead of the
+default per-cpu set.  This can be useful for applications interested
+mainly in seeing the relative ordering of system-wide events without
+the need to bother with saving explicit timestamps for the purpose of
+merging/sorting per-cpu files in a postprocessing step.
+
+To have relay_open() create a global buffer, the create_buf_file()
+implementation should set the value of the is_global outparam to a
+non-zero value in addition to creating the file that will be used to
+represent the single buffer.  In the case of a global buffer,
+create_buf_file() and remove_buf_file() will be called only once.  The
+normal channel-writing functions, e.g. relay_write(), can still be
+used - writes from any cpu will transparently end up in the global
+buffer - but since it is a global buffer, callers should make sure
+they use the proper locking for such a buffer, either by wrapping
+writes in a spinlock, or by copying a write function from relay.h and
+creating a local version that internally does the proper locking.
+
+The private_data passed into relay_open() allows clients to associate
+user-defined data with a channel, and is immediately available
+(including in create_buf_file()) via chan->private_data or
+buf->chan->private_data.
+
+Buffer-only channels
+--------------------
+
+These channels have no files associated and can be created with
+relay_open(NULL, NULL, ...). Such channels are useful in scenarios such
+as when doing early tracing in the kernel, before the VFS is up. In these
+cases, one may open a buffer-only channel and then call
+relay_late_setup_files() when the kernel is ready to handle files,
+to expose the buffered data to the userspace.
+
+Channel 'modes'
+---------------
+
+relay channels can be used in either of two modes - 'overwrite' or
+'no-overwrite'.  The mode is entirely determined by the implementation
+of the subbuf_start() callback, as described below.  The default if no
+subbuf_start() callback is defined is 'no-overwrite' mode.  If the
+default mode suits your needs, and you plan to use the read()
+interface to retrieve channel data, you can ignore the details of this
+section, as it pertains mainly to mmap() implementations.
+
+In 'overwrite' mode, also known as 'flight recorder' mode, writes
+continuously cycle around the buffer and will never fail, but will
+unconditionally overwrite old data regardless of whether it's actually
+been consumed.  In no-overwrite mode, writes will fail, i.e. data will
+be lost, if the number of unconsumed sub-buffers equals the total
+number of sub-buffers in the channel.  It should be clear that if
+there is no consumer or if the consumer can't consume sub-buffers fast
+enough, data will be lost in either case; the only difference is
+whether data is lost from the beginning or the end of a buffer.
+
+As explained above, a relay channel is made of up one or more
+per-cpu channel buffers, each implemented as a circular buffer
+subdivided into one or more sub-buffers.  Messages are written into
+the current sub-buffer of the channel's current per-cpu buffer via the
+write functions described below.  Whenever a message can't fit into
+the current sub-buffer, because there's no room left for it, the
+client is notified via the subbuf_start() callback that a switch to a
+new sub-buffer is about to occur.  The client uses this callback to 1)
+initialize the next sub-buffer if appropriate 2) finalize the previous
+sub-buffer if appropriate and 3) return a boolean value indicating
+whether or not to actually move on to the next sub-buffer.
+
+To implement 'no-overwrite' mode, the userspace client would provide
+an implementation of the subbuf_start() callback something like the
+following:
+
+static int subbuf_start(struct rchan_buf *buf,
+                        void *subbuf,
+			void *prev_subbuf,
+			unsigned int prev_padding)
+{
+	if (prev_subbuf)
+		*((unsigned *)prev_subbuf) = prev_padding;
+
+	if (relay_buf_full(buf))
+		return 0;
+
+	subbuf_start_reserve(buf, sizeof(unsigned int));
+
+	return 1;
+}
+
+If the current buffer is full, i.e. all sub-buffers remain unconsumed,
+the callback returns 0 to indicate that the buffer switch should not
+occur yet, i.e. until the consumer has had a chance to read the
+current set of ready sub-buffers.  For the relay_buf_full() function
+to make sense, the consumer is responsible for notifying the relay
+interface when sub-buffers have been consumed via
+relay_subbufs_consumed().  Any subsequent attempts to write into the
+buffer will again invoke the subbuf_start() callback with the same
+parameters; only when the consumer has consumed one or more of the
+ready sub-buffers will relay_buf_full() return 0, in which case the
+buffer switch can continue.
+
+The implementation of the subbuf_start() callback for 'overwrite' mode
+would be very similar:
+
+static int subbuf_start(struct rchan_buf *buf,
+                        void *subbuf,
+			void *prev_subbuf,
+			size_t prev_padding)
+{
+	if (prev_subbuf)
+		*((unsigned *)prev_subbuf) = prev_padding;
+
+	subbuf_start_reserve(buf, sizeof(unsigned int));
+
+	return 1;
+}
+
+In this case, the relay_buf_full() check is meaningless and the
+callback always returns 1, causing the buffer switch to occur
+unconditionally.  It's also meaningless for the client to use the
+relay_subbufs_consumed() function in this mode, as it's never
+consulted.
+
+The default subbuf_start() implementation, used if the client doesn't
+define any callbacks, or doesn't define the subbuf_start() callback,
+implements the simplest possible 'no-overwrite' mode, i.e. it does
+nothing but return 0.
+
+Header information can be reserved at the beginning of each sub-buffer
+by calling the subbuf_start_reserve() helper function from within the
+subbuf_start() callback.  This reserved area can be used to store
+whatever information the client wants.  In the example above, room is
+reserved in each sub-buffer to store the padding count for that
+sub-buffer.  This is filled in for the previous sub-buffer in the
+subbuf_start() implementation; the padding value for the previous
+sub-buffer is passed into the subbuf_start() callback along with a
+pointer to the previous sub-buffer, since the padding value isn't
+known until a sub-buffer is filled.  The subbuf_start() callback is
+also called for the first sub-buffer when the channel is opened, to
+give the client a chance to reserve space in it.  In this case the
+previous sub-buffer pointer passed into the callback will be NULL, so
+the client should check the value of the prev_subbuf pointer before
+writing into the previous sub-buffer.
+
+Writing to a channel
+--------------------
+
+Kernel clients write data into the current cpu's channel buffer using
+relay_write() or __relay_write().  relay_write() is the main logging
+function - it uses local_irqsave() to protect the buffer and should be
+used if you might be logging from interrupt context.  If you know
+you'll never be logging from interrupt context, you can use
+__relay_write(), which only disables preemption.  These functions
+don't return a value, so you can't determine whether or not they
+failed - the assumption is that you wouldn't want to check a return
+value in the fast logging path anyway, and that they'll always succeed
+unless the buffer is full and no-overwrite mode is being used, in
+which case you can detect a failed write in the subbuf_start()
+callback by calling the relay_buf_full() helper function.
+
+relay_reserve() is used to reserve a slot in a channel buffer which
+can be written to later.  This would typically be used in applications
+that need to write directly into a channel buffer without having to
+stage data in a temporary buffer beforehand.  Because the actual write
+may not happen immediately after the slot is reserved, applications
+using relay_reserve() can keep a count of the number of bytes actually
+written, either in space reserved in the sub-buffers themselves or as
+a separate array.  See the 'reserve' example in the relay-apps tarball
+at http://relayfs.sourceforge.net for an example of how this can be
+done.  Because the write is under control of the client and is
+separated from the reserve, relay_reserve() doesn't protect the buffer
+at all - it's up to the client to provide the appropriate
+synchronization when using relay_reserve().
+
+Closing a channel
+-----------------
+
+The client calls relay_close() when it's finished using the channel.
+The channel and its associated buffers are destroyed when there are no
+longer any references to any of the channel buffers.  relay_flush()
+forces a sub-buffer switch on all the channel buffers, and can be used
+to finalize and process the last sub-buffers before the channel is
+closed.
+
+Misc
+----
+
+Some applications may want to keep a channel around and re-use it
+rather than open and close a new channel for each use.  relay_reset()
+can be used for this purpose - it resets a channel to its initial
+state without reallocating channel buffer memory or destroying
+existing mappings.  It should however only be called when it's safe to
+do so, i.e. when the channel isn't currently being written to.
+
+Finally, there are a couple of utility callbacks that can be used for
+different purposes.  buf_mapped() is called whenever a channel buffer
+is mmapped from user space and buf_unmapped() is called when it's
+unmapped.  The client can use this notification to trigger actions
+within the kernel application, such as enabling/disabling logging to
+the channel.
+
+
+Resources
+=========
+
+For news, example code, mailing list, etc. see the relay interface homepage:
+
+    http://relayfs.sourceforge.net
+
+
+Credits
+=======
+
+The ideas and specs for the relay interface came about as a result of
+discussions on tracing involving the following:
+
+Michel Dagenais		<michel.dagenais@polymtl.ca>
+Richard Moore		<richardj_moore@uk.ibm.com>
+Bob Wisniewski		<bob@watson.ibm.com>
+Karim Yaghmour		<karim@opersys.com>
+Tom Zanussi		<zanussi@us.ibm.com>
+
+Also thanks to Hubertus Franke for a lot of useful suggestions and bug
+reports.
diff --git a/Documentation/filesystems/romfs.txt b/Documentation/filesystems/romfs.txt
new file mode 100644
index 000000000..e2b07cc91
--- /dev/null
+++ b/Documentation/filesystems/romfs.txt
@@ -0,0 +1,186 @@
+ROMFS - ROM FILE SYSTEM
+
+This is a quite dumb, read only filesystem, mainly for initial RAM
+disks of installation disks.  It has grown up by the need of having
+modules linked at boot time.  Using this filesystem, you get a very
+similar feature, and even the possibility of a small kernel, with a
+file system which doesn't take up useful memory from the router
+functions in the basement of your office.
+
+For comparison, both the older minix and xiafs (the latter is now
+defunct) filesystems, compiled as module need more than 20000 bytes,
+while romfs is less than a page, about 4000 bytes (assuming i586
+code).  Under the same conditions, the msdos filesystem would need
+about 30K (and does not support device nodes or symlinks), while the
+nfs module with nfsroot is about 57K.  Furthermore, as a bit unfair
+comparison, an actual rescue disk used up 3202 blocks with ext2, while
+with romfs, it needed 3079 blocks.
+
+To create such a file system, you'll need a user program named
+genromfs. It is available on http://romfs.sourceforge.net/
+
+As the name suggests, romfs could be also used (space-efficiently) on
+various read-only media, like (E)EPROM disks if someone will have the
+motivation.. :)
+
+However, the main purpose of romfs is to have a very small kernel,
+which has only this filesystem linked in, and then can load any module
+later, with the current module utilities.  It can also be used to run
+some program to decide if you need SCSI devices, and even IDE or
+floppy drives can be loaded later if you use the "initrd"--initial
+RAM disk--feature of the kernel.  This would not be really news
+flash, but with romfs, you can even spare off your ext2 or minix or
+maybe even affs filesystem until you really know that you need it.
+
+For example, a distribution boot disk can contain only the cd disk
+drivers (and possibly the SCSI drivers), and the ISO 9660 filesystem
+module.  The kernel can be small enough, since it doesn't have other
+filesystems, like the quite large ext2fs module, which can then be
+loaded off the CD at a later stage of the installation.  Another use
+would be for a recovery disk, when you are reinstalling a workstation
+from the network, and you will have all the tools/modules available
+from a nearby server, so you don't want to carry two disks for this
+purpose, just because it won't fit into ext2.
+
+romfs operates on block devices as you can expect, and the underlying
+structure is very simple.  Every accessible structure begins on 16
+byte boundaries for fast access.  The minimum space a file will take
+is 32 bytes (this is an empty file, with a less than 16 character
+name).  The maximum overhead for any non-empty file is the header, and
+the 16 byte padding for the name and the contents, also 16+14+15 = 45
+bytes.  This is quite rare however, since most file names are longer
+than 3 bytes, and shorter than 15 bytes.
+
+The layout of the filesystem is the following:
+
+offset	    content
+
+	+---+---+---+---+
+  0	| - | r | o | m |  \
+	+---+---+---+---+	The ASCII representation of those bytes
+  4	| 1 | f | s | - |  /	(i.e. "-rom1fs-")
+	+---+---+---+---+
+  8	|   full size	|	The number of accessible bytes in this fs.
+	+---+---+---+---+
+ 12	|    checksum	|	The checksum of the FIRST 512 BYTES.
+	+---+---+---+---+
+ 16	| volume name	|	The zero terminated name of the volume,
+	:               :	padded to 16 byte boundary.
+	+---+---+---+---+
+ xx	|     file	|
+	:    headers	:
+
+Every multi byte value (32 bit words, I'll use the longwords term from
+now on) must be in big endian order.
+
+The first eight bytes identify the filesystem, even for the casual
+inspector.  After that, in the 3rd longword, it contains the number of
+bytes accessible from the start of this filesystem.  The 4th longword
+is the checksum of the first 512 bytes (or the number of bytes
+accessible, whichever is smaller).  The applied algorithm is the same
+as in the AFFS filesystem, namely a simple sum of the longwords
+(assuming bigendian quantities again).  For details, please consult
+the source.  This algorithm was chosen because although it's not quite
+reliable, it does not require any tables, and it is very simple.
+
+The following bytes are now part of the file system; each file header
+must begin on a 16 byte boundary.
+
+offset	    content
+
+     	+---+---+---+---+
+  0	| next filehdr|X|	The offset of the next file header
+	+---+---+---+---+	  (zero if no more files)
+  4	|   spec.info	|	Info for directories/hard links/devices
+	+---+---+---+---+
+  8	|     size      |	The size of this file in bytes
+	+---+---+---+---+
+ 12	|   checksum	|	Covering the meta data, including the file
+	+---+---+---+---+	  name, and padding
+ 16	| file name     |	The zero terminated name of the file,
+	:               :	padded to 16 byte boundary
+	+---+---+---+---+
+ xx	| file data	|
+	:		:
+
+Since the file headers begin always at a 16 byte boundary, the lowest
+4 bits would be always zero in the next filehdr pointer.  These four
+bits are used for the mode information.  Bits 0..2 specify the type of
+the file; while bit 4 shows if the file is executable or not.  The
+permissions are assumed to be world readable, if this bit is not set,
+and world executable if it is; except the character and block devices,
+they are never accessible for other than owner.  The owner of every
+file is user and group 0, this should never be a problem for the
+intended use.  The mapping of the 8 possible values to file types is
+the following:
+
+	  mapping		spec.info means
+ 0	hard link	link destination [file header]
+ 1	directory	first file's header
+ 2	regular file	unused, must be zero [MBZ]
+ 3	symbolic link	unused, MBZ (file data is the link content)
+ 4	block device	16/16 bits major/minor number
+ 5	char device		    - " -
+ 6	socket		unused, MBZ
+ 7	fifo		unused, MBZ
+
+Note that hard links are specifically marked in this filesystem, but
+they will behave as you can expect (i.e. share the inode number).
+Note also that it is your responsibility to not create hard link
+loops, and creating all the . and .. links for directories.  This is
+normally done correctly by the genromfs program.  Please refrain from
+using the executable bits for special purposes on the socket and fifo
+special files, they may have other uses in the future.  Additionally,
+please remember that only regular files, and symlinks are supposed to
+have a nonzero size field; they contain the number of bytes available
+directly after the (padded) file name.
+
+Another thing to note is that romfs works on file headers and data
+aligned to 16 byte boundaries, but most hardware devices and the block
+device drivers are unable to cope with smaller than block-sized data.
+To overcome this limitation, the whole size of the file system must be
+padded to an 1024 byte boundary.
+
+If you have any problems or suggestions concerning this file system,
+please contact me.  However, think twice before wanting me to add
+features and code, because the primary and most important advantage of
+this file system is the small code.  On the other hand, don't be
+alarmed, I'm not getting that much romfs related mail.  Now I can
+understand why Avery wrote poems in the ARCnet docs to get some more
+feedback. :)
+
+romfs has also a mailing list, and to date, it hasn't received any
+traffic, so you are welcome to join it to discuss your ideas. :)
+
+It's run by ezmlm, so you can subscribe to it by sending a message
+to romfs-subscribe@shadow.banki.hu, the content is irrelevant.
+
+Pending issues:
+
+- Permissions and owner information are pretty essential features of a
+Un*x like system, but romfs does not provide the full possibilities.
+I have never found this limiting, but others might.
+
+- The file system is read only, so it can be very small, but in case
+one would want to write _anything_ to a file system, he still needs
+a writable file system, thus negating the size advantages.  Possible
+solutions: implement write access as a compile-time option, or a new,
+similarly small writable filesystem for RAM disks.
+
+- Since the files are only required to have alignment on a 16 byte
+boundary, it is currently possibly suboptimal to read or execute files
+from the filesystem.  It might be resolved by reordering file data to
+have most of it (i.e. except the start and the end) laying at "natural"
+boundaries, thus it would be possible to directly map a big portion of
+the file contents to the mm subsystem.
+
+- Compression might be an useful feature, but memory is quite a
+limiting factor in my eyes.
+
+- Where it is used?
+
+- Does it work on other architectures than intel and motorola?
+
+
+Have fun,
+Janos Farkas <chexum@shadow.banki.hu>
diff --git a/Documentation/filesystems/seq_file.txt b/Documentation/filesystems/seq_file.txt
new file mode 100644
index 000000000..7cf714392
--- /dev/null
+++ b/Documentation/filesystems/seq_file.txt
@@ -0,0 +1,365 @@
+The seq_file interface
+
+	Copyright 2003 Jonathan Corbet <corbet@lwn.net>
+	This file is originally from the LWN.net Driver Porting series at
+	http://lwn.net/Articles/driver-porting/
+
+
+There are numerous ways for a device driver (or other kernel component) to
+provide information to the user or system administrator.  One useful
+technique is the creation of virtual files, in debugfs, /proc or elsewhere.
+Virtual files can provide human-readable output that is easy to get at
+without any special utility programs; they can also make life easier for
+script writers. It is not surprising that the use of virtual files has
+grown over the years.
+
+Creating those files correctly has always been a bit of a challenge,
+however. It is not that hard to make a virtual file which returns a
+string. But life gets trickier if the output is long - anything greater
+than an application is likely to read in a single operation.  Handling
+multiple reads (and seeks) requires careful attention to the reader's
+position within the virtual file - that position is, likely as not, in the
+middle of a line of output. The kernel has traditionally had a number of
+implementations that got this wrong.
+
+The 2.6 kernel contains a set of functions (implemented by Alexander Viro)
+which are designed to make it easy for virtual file creators to get it
+right.
+
+The seq_file interface is available via <linux/seq_file.h>. There are
+three aspects to seq_file:
+
+     * An iterator interface which lets a virtual file implementation
+       step through the objects it is presenting.
+
+     * Some utility functions for formatting objects for output without
+       needing to worry about things like output buffers.
+
+     * A set of canned file_operations which implement most operations on
+       the virtual file.
+
+We'll look at the seq_file interface via an extremely simple example: a
+loadable module which creates a file called /proc/sequence. The file, when
+read, simply produces a set of increasing integer values, one per line. The
+sequence will continue until the user loses patience and finds something
+better to do. The file is seekable, in that one can do something like the
+following:
+
+    dd if=/proc/sequence of=out1 count=1
+    dd if=/proc/sequence skip=1 of=out2 count=1
+
+Then concatenate the output files out1 and out2 and get the right
+result. Yes, it is a thoroughly useless module, but the point is to show
+how the mechanism works without getting lost in other details.  (Those
+wanting to see the full source for this module can find it at
+http://lwn.net/Articles/22359/).
+
+Deprecated create_proc_entry
+
+Note that the above article uses create_proc_entry which was removed in
+kernel 3.10. Current versions require the following update
+
+-	entry = create_proc_entry("sequence", 0, NULL);
+-	if (entry)
+-		entry->proc_fops = &ct_file_ops;
++	entry = proc_create("sequence", 0, NULL, &ct_file_ops);
+
+The iterator interface
+
+Modules implementing a virtual file with seq_file must implement an
+iterator object that allows stepping through the data of interest
+during a "session" (roughly one read() system call).  If the iterator
+is able to move to a specific position - like the file they implement,
+though with freedom to map the position number to a sequence location
+in whatever way is convenient - the iterator need only exist
+transiently during a session.  If the iterator cannot easily find a
+numerical position but works well with a first/next interface, the
+iterator can be stored in the private data area and continue from one
+session to the next.
+
+A seq_file implementation that is formatting firewall rules from a
+table, for example, could provide a simple iterator that interprets
+position N as the Nth rule in the chain.  A seq_file implementation
+that presents the content of a, potentially volatile, linked list
+might record a pointer into that list, providing that can be done
+without risk of the current location being removed.
+
+Positioning can thus be done in whatever way makes the most sense for
+the generator of the data, which need not be aware of how a position
+translates to an offset in the virtual file. The one obvious exception
+is that a position of zero should indicate the beginning of the file.
+
+The /proc/sequence iterator just uses the count of the next number it
+will output as its position.
+
+Four functions must be implemented to make the iterator work. The
+first, called start(), starts a session and takes a position as an
+argument, returning an iterator which will start reading at that
+position.  The pos passed to start() will always be either zero, or
+the most recent pos used in the previous session.
+
+For our simple sequence example,
+the start() function looks like:
+
+	static void *ct_seq_start(struct seq_file *s, loff_t *pos)
+	{
+	        loff_t *spos = kmalloc(sizeof(loff_t), GFP_KERNEL);
+	        if (! spos)
+	                return NULL;
+	        *spos = *pos;
+	        return spos;
+	}
+
+The entire data structure for this iterator is a single loff_t value
+holding the current position. There is no upper bound for the sequence
+iterator, but that will not be the case for most other seq_file
+implementations; in most cases the start() function should check for a
+"past end of file" condition and return NULL if need be.
+
+For more complicated applications, the private field of the seq_file
+structure can be used to hold state from session to session.  There is
+also a special value which can be returned by the start() function
+called SEQ_START_TOKEN; it can be used if you wish to instruct your
+show() function (described below) to print a header at the top of the
+output. SEQ_START_TOKEN should only be used if the offset is zero,
+however.
+
+The next function to implement is called, amazingly, next(); its job is to
+move the iterator forward to the next position in the sequence.  The
+example module can simply increment the position by one; more useful
+modules will do what is needed to step through some data structure. The
+next() function returns a new iterator, or NULL if the sequence is
+complete. Here's the example version:
+
+	static void *ct_seq_next(struct seq_file *s, void *v, loff_t *pos)
+	{
+	        loff_t *spos = v;
+	        *pos = ++*spos;
+	        return spos;
+	}
+
+The stop() function closes a session; its job, of course, is to clean
+up. If dynamic memory is allocated for the iterator, stop() is the
+place to free it; if a lock was taken by start(), stop() must release
+that lock.  The value that *pos was set to by the last next() call
+before stop() is remembered, and used for the first start() call of
+the next session unless lseek() has been called on the file; in that
+case next start() will be asked to start at position zero.
+
+	static void ct_seq_stop(struct seq_file *s, void *v)
+	{
+	        kfree(v);
+	}
+
+Finally, the show() function should format the object currently pointed to
+by the iterator for output.  The example module's show() function is:
+
+	static int ct_seq_show(struct seq_file *s, void *v)
+	{
+	        loff_t *spos = v;
+	        seq_printf(s, "%lld\n", (long long)*spos);
+	        return 0;
+	}
+
+If all is well, the show() function should return zero.  A negative error
+code in the usual manner indicates that something went wrong; it will be
+passed back to user space.  This function can also return SEQ_SKIP, which
+causes the current item to be skipped; if the show() function has already
+generated output before returning SEQ_SKIP, that output will be dropped.
+
+We will look at seq_printf() in a moment. But first, the definition of the
+seq_file iterator is finished by creating a seq_operations structure with
+the four functions we have just defined:
+
+	static const struct seq_operations ct_seq_ops = {
+	        .start = ct_seq_start,
+	        .next  = ct_seq_next,
+	        .stop  = ct_seq_stop,
+	        .show  = ct_seq_show
+	};
+
+This structure will be needed to tie our iterator to the /proc file in
+a little bit.
+
+It's worth noting that the iterator value returned by start() and
+manipulated by the other functions is considered to be completely opaque by
+the seq_file code. It can thus be anything that is useful in stepping
+through the data to be output. Counters can be useful, but it could also be
+a direct pointer into an array or linked list. Anything goes, as long as
+the programmer is aware that things can happen between calls to the
+iterator function. However, the seq_file code (by design) will not sleep
+between the calls to start() and stop(), so holding a lock during that time
+is a reasonable thing to do. The seq_file code will also avoid taking any
+other locks while the iterator is active.
+
+The iterater value returned by start() or next() is guaranteed to be
+passed to a subsequent next() or stop() call.  This allows resources
+such as locks that were taken to be reliably released.  There is *no*
+guarantee that the iterator will be passed to show(), though in practice
+it often will be.
+
+
+Formatted output
+
+The seq_file code manages positioning within the output created by the
+iterator and getting it into the user's buffer. But, for that to work, that
+output must be passed to the seq_file code. Some utility functions have
+been defined which make this task easy.
+
+Most code will simply use seq_printf(), which works pretty much like
+printk(), but which requires the seq_file pointer as an argument.
+
+For straight character output, the following functions may be used:
+
+	seq_putc(struct seq_file *m, char c);
+	seq_puts(struct seq_file *m, const char *s);
+	seq_escape(struct seq_file *m, const char *s, const char *esc);
+
+The first two output a single character and a string, just like one would
+expect. seq_escape() is like seq_puts(), except that any character in s
+which is in the string esc will be represented in octal form in the output.
+
+There are also a pair of functions for printing filenames:
+
+	int seq_path(struct seq_file *m, const struct path *path,
+		     const char *esc);
+	int seq_path_root(struct seq_file *m, const struct path *path,
+			  const struct path *root, const char *esc)
+
+Here, path indicates the file of interest, and esc is a set of characters
+which should be escaped in the output.  A call to seq_path() will output
+the path relative to the current process's filesystem root.  If a different
+root is desired, it can be used with seq_path_root().  If it turns out that
+path cannot be reached from root, seq_path_root() returns SEQ_SKIP.
+
+A function producing complicated output may want to check
+	bool seq_has_overflowed(struct seq_file *m);
+and avoid further seq_<output> calls if true is returned.
+
+A true return from seq_has_overflowed means that the seq_file buffer will
+be discarded and the seq_show function will attempt to allocate a larger
+buffer and retry printing.
+
+
+Making it all work
+
+So far, we have a nice set of functions which can produce output within the
+seq_file system, but we have not yet turned them into a file that a user
+can see. Creating a file within the kernel requires, of course, the
+creation of a set of file_operations which implement the operations on that
+file. The seq_file interface provides a set of canned operations which do
+most of the work. The virtual file author still must implement the open()
+method, however, to hook everything up. The open function is often a single
+line, as in the example module:
+
+	static int ct_open(struct inode *inode, struct file *file)
+	{
+		return seq_open(file, &ct_seq_ops);
+	}
+
+Here, the call to seq_open() takes the seq_operations structure we created
+before, and gets set up to iterate through the virtual file.
+
+On a successful open, seq_open() stores the struct seq_file pointer in
+file->private_data. If you have an application where the same iterator can
+be used for more than one file, you can store an arbitrary pointer in the
+private field of the seq_file structure; that value can then be retrieved
+by the iterator functions.
+
+There is also a wrapper function to seq_open() called seq_open_private(). It
+kmallocs a zero filled block of memory and stores a pointer to it in the
+private field of the seq_file structure, returning 0 on success. The
+block size is specified in a third parameter to the function, e.g.:
+
+	static int ct_open(struct inode *inode, struct file *file)
+	{
+		return seq_open_private(file, &ct_seq_ops,
+					sizeof(struct mystruct));
+	}
+
+There is also a variant function, __seq_open_private(), which is functionally
+identical except that, if successful, it returns the pointer to the allocated
+memory block, allowing further initialisation e.g.:
+
+	static int ct_open(struct inode *inode, struct file *file)
+	{
+		struct mystruct *p =
+			__seq_open_private(file, &ct_seq_ops, sizeof(*p));
+
+		if (!p)
+			return -ENOMEM;
+
+		p->foo = bar; /* initialize my stuff */
+			...
+		p->baz = true;
+
+		return 0;
+	}
+
+A corresponding close function, seq_release_private() is available which
+frees the memory allocated in the corresponding open.
+
+The other operations of interest - read(), llseek(), and release() - are
+all implemented by the seq_file code itself. So a virtual file's
+file_operations structure will look like:
+
+	static const struct file_operations ct_file_ops = {
+	        .owner   = THIS_MODULE,
+	        .open    = ct_open,
+	        .read    = seq_read,
+	        .llseek  = seq_lseek,
+	        .release = seq_release
+	};
+
+There is also a seq_release_private() which passes the contents of the
+seq_file private field to kfree() before releasing the structure.
+
+The final step is the creation of the /proc file itself. In the example
+code, that is done in the initialization code in the usual way:
+
+	static int ct_init(void)
+	{
+	        struct proc_dir_entry *entry;
+
+	        proc_create("sequence", 0, NULL, &ct_file_ops);
+	        return 0;
+	}
+
+	module_init(ct_init);
+
+And that is pretty much it.
+
+
+seq_list
+
+If your file will be iterating through a linked list, you may find these
+routines useful:
+
+	struct list_head *seq_list_start(struct list_head *head,
+	       		 		 loff_t pos);
+	struct list_head *seq_list_start_head(struct list_head *head,
+			 		      loff_t pos);
+	struct list_head *seq_list_next(void *v, struct list_head *head,
+					loff_t *ppos);
+
+These helpers will interpret pos as a position within the list and iterate
+accordingly.  Your start() and next() functions need only invoke the
+seq_list_* helpers with a pointer to the appropriate list_head structure.
+
+
+The extra-simple version
+
+For extremely simple virtual files, there is an even easier interface.  A
+module can define only the show() function, which should create all the
+output that the virtual file will contain. The file's open() method then
+calls:
+
+	int single_open(struct file *file,
+	                int (*show)(struct seq_file *m, void *p),
+	                void *data);
+
+When output time comes, the show() function will be called once. The data
+value given to single_open() can be found in the private field of the
+seq_file structure. When using single_open(), the programmer should use
+single_release() instead of seq_release() in the file_operations structure
+to avoid a memory leak.
diff --git a/Documentation/filesystems/sharedsubtree.txt b/Documentation/filesystems/sharedsubtree.txt
new file mode 100644
index 000000000..8ccfbd552
--- /dev/null
+++ b/Documentation/filesystems/sharedsubtree.txt
@@ -0,0 +1,939 @@
+Shared Subtrees
+---------------
+
+Contents:
+	1) Overview
+	2) Features
+	3) Setting mount states
+	4) Use-case
+	5) Detailed semantics
+	6) Quiz
+	7) FAQ
+	8) Implementation
+
+
+1) Overview
+-----------
+
+Consider the following situation:
+
+A process wants to clone its own namespace, but still wants to access the CD
+that got mounted recently.  Shared subtree semantics provide the necessary
+mechanism to accomplish the above.
+
+It provides the necessary building blocks for features like per-user-namespace
+and versioned filesystem.
+
+2) Features
+-----------
+
+Shared subtree provides four different flavors of mounts; struct vfsmount to be
+precise
+
+	a. shared mount
+	b. slave mount
+	c. private mount
+	d. unbindable mount
+
+
+2a) A shared mount can be replicated to as many mountpoints and all the
+replicas continue to be exactly same.
+
+	Here is an example:
+
+	Let's say /mnt has a mount that is shared.
+	mount --make-shared /mnt
+
+	Note: mount(8) command now supports the --make-shared flag,
+	so the sample 'smount' program is no longer needed and has been
+	removed.
+
+	# mount --bind /mnt /tmp
+	The above command replicates the mount at /mnt to the mountpoint /tmp
+	and the contents of both the mounts remain identical.
+
+	#ls /mnt
+	a b c
+
+	#ls /tmp
+	a b c
+
+	Now let's say we mount a device at /tmp/a
+	# mount /dev/sd0  /tmp/a
+
+	#ls /tmp/a
+	t1 t2 t3
+
+	#ls /mnt/a
+	t1 t2 t3
+
+	Note that the mount has propagated to the mount at /mnt as well.
+
+	And the same is true even when /dev/sd0 is mounted on /mnt/a. The
+	contents will be visible under /tmp/a too.
+
+
+2b) A slave mount is like a shared mount except that mount and umount events
+	only propagate towards it.
+
+	All slave mounts have a master mount which is a shared.
+
+	Here is an example:
+
+	Let's say /mnt has a mount which is shared.
+	# mount --make-shared /mnt
+
+	Let's bind mount /mnt to /tmp
+	# mount --bind /mnt /tmp
+
+	the new mount at /tmp becomes a shared mount and it is a replica of
+	the mount at /mnt.
+
+	Now let's make the mount at /tmp; a slave of /mnt
+	# mount --make-slave /tmp
+
+	let's mount /dev/sd0 on /mnt/a
+	# mount /dev/sd0 /mnt/a
+
+	#ls /mnt/a
+	t1 t2 t3
+
+	#ls /tmp/a
+	t1 t2 t3
+
+	Note the mount event has propagated to the mount at /tmp
+
+	However let's see what happens if we mount something on the mount at /tmp
+
+	# mount /dev/sd1 /tmp/b
+
+	#ls /tmp/b
+	s1 s2 s3
+
+	#ls /mnt/b
+
+	Note how the mount event has not propagated to the mount at
+	/mnt
+
+
+2c) A private mount does not forward or receive propagation.
+
+	This is the mount we are familiar with. Its the default type.
+
+
+2d) A unbindable mount is a unbindable private mount
+
+	let's say we have a mount at /mnt and we make it unbindable
+
+	# mount --make-unbindable /mnt
+
+	 Let's try to bind mount this mount somewhere else.
+	 # mount --bind /mnt /tmp
+	 mount: wrong fs type, bad option, bad superblock on /mnt,
+	        or too many mounted file systems
+
+	Binding a unbindable mount is a invalid operation.
+
+
+3) Setting mount states
+
+	The mount command (util-linux package) can be used to set mount
+	states:
+
+	mount --make-shared mountpoint
+	mount --make-slave mountpoint
+	mount --make-private mountpoint
+	mount --make-unbindable mountpoint
+
+
+4) Use cases
+------------
+
+	A) A process wants to clone its own namespace, but still wants to
+	   access the CD that got mounted recently.
+
+	   Solution:
+
+		The system administrator can make the mount at /cdrom shared
+		mount --bind /cdrom /cdrom
+		mount --make-shared /cdrom
+
+		Now any process that clones off a new namespace will have a
+		mount at /cdrom which is a replica of the same mount in the
+		parent namespace.
+
+		So when a CD is inserted and mounted at /cdrom that mount gets
+		propagated to the other mount at /cdrom in all the other clone
+		namespaces.
+
+	B) A process wants its mounts invisible to any other process, but
+	still be able to see the other system mounts.
+
+	   Solution:
+
+		To begin with, the administrator can mark the entire mount tree
+		as shareable.
+
+		mount --make-rshared /
+
+		A new process can clone off a new namespace. And mark some part
+		of its namespace as slave
+
+		mount --make-rslave /myprivatetree
+
+		Hence forth any mounts within the /myprivatetree done by the
+		process will not show up in any other namespace. However mounts
+		done in the parent namespace under /myprivatetree still shows
+		up in the process's namespace.
+
+
+	Apart from the above semantics this feature provides the
+	building blocks to solve the following problems:
+
+	C)  Per-user namespace
+
+		The above semantics allows a way to share mounts across
+		namespaces.  But namespaces are associated with processes. If
+		namespaces are made first class objects with user API to
+		associate/disassociate a namespace with userid, then each user
+		could have his/her own namespace and tailor it to his/her
+		requirements. This needs to be supported in PAM.
+
+	D)  Versioned files
+
+		If the entire mount tree is visible at multiple locations, then
+		an underlying versioning file system can return different
+		versions of the file depending on the path used to access that
+		file.
+
+		An example is:
+
+		mount --make-shared /
+		mount --rbind / /view/v1
+		mount --rbind / /view/v2
+		mount --rbind / /view/v3
+		mount --rbind / /view/v4
+
+		and if /usr has a versioning filesystem mounted, then that
+		mount appears at /view/v1/usr, /view/v2/usr, /view/v3/usr and
+		/view/v4/usr too
+
+		A user can request v3 version of the file /usr/fs/namespace.c
+		by accessing /view/v3/usr/fs/namespace.c . The underlying
+		versioning filesystem can then decipher that v3 version of the
+		filesystem is being requested and return the corresponding
+		inode.
+
+5) Detailed semantics:
+-------------------
+	The section below explains the detailed semantics of
+	bind, rbind, move, mount, umount and clone-namespace operations.
+
+	Note: the word 'vfsmount' and the noun 'mount' have been used
+	to mean the same thing, throughout this document.
+
+5a) Mount states
+
+	A given mount can be in one of the following states
+	1) shared
+	2) slave
+	3) shared and slave
+	4) private
+	5) unbindable
+
+	A 'propagation event' is defined as event generated on a vfsmount
+	that leads to mount or unmount actions in other vfsmounts.
+
+	A 'peer group' is defined as a group of vfsmounts that propagate
+	events to each other.
+
+	(1) Shared mounts
+
+		A 'shared mount' is defined as a vfsmount that belongs to a
+		'peer group'.
+
+		For example:
+			mount --make-shared /mnt
+			mount --bind /mnt /tmp
+
+		The mount at /mnt and that at /tmp are both shared and belong
+		to the same peer group. Anything mounted or unmounted under
+		/mnt or /tmp reflect in all the other mounts of its peer
+		group.
+
+
+	(2) Slave mounts
+
+		A 'slave mount' is defined as a vfsmount that receives
+		propagation events and does not forward propagation events.
+
+		A slave mount as the name implies has a master mount from which
+		mount/unmount events are received. Events do not propagate from
+		the slave mount to the master.  Only a shared mount can be made
+		a slave by executing the following command
+
+			mount --make-slave mount
+
+		A shared mount that is made as a slave is no more shared unless
+		modified to become shared.
+
+	(3) Shared and Slave
+
+		A vfsmount can be both shared as well as slave.  This state
+		indicates that the mount is a slave of some vfsmount, and
+		has its own peer group too.  This vfsmount receives propagation
+		events from its master vfsmount, and also forwards propagation
+		events to its 'peer group' and to its slave vfsmounts.
+
+		Strictly speaking, the vfsmount is shared having its own
+		peer group, and this peer-group is a slave of some other
+		peer group.
+
+		Only a slave vfsmount can be made as 'shared and slave' by
+		either executing the following command
+			mount --make-shared mount
+		or by moving the slave vfsmount under a shared vfsmount.
+
+	(4) Private mount
+
+		A 'private mount' is defined as vfsmount that does not
+		receive or forward any propagation events.
+
+	(5) Unbindable mount
+
+		A 'unbindable mount' is defined as vfsmount that does not
+		receive or forward any propagation events and cannot
+		be bind mounted.
+
+
+   	State diagram:
+   	The state diagram below explains the state transition of a mount,
+	in response to various commands.
+	------------------------------------------------------------------------
+	|             |make-shared |  make-slave  | make-private |make-unbindab|
+	--------------|------------|--------------|--------------|-------------|
+	|shared	      |shared	   |*slave/private|   private	 | unbindable  |
+	|             |            |              |              |             |
+	|-------------|------------|--------------|--------------|-------------|
+	|slave	      |shared      |	**slave	  |    private   | unbindable  |
+	|             |and slave   |              |              |             |
+	|-------------|------------|--------------|--------------|-------------|
+	|shared	      |shared      |    slave	  |    private   | unbindable  |
+	|and slave    |and slave   |              |              |             |
+	|-------------|------------|--------------|--------------|-------------|
+	|private      |shared	   |  **private	  |    private   | unbindable  |
+	|-------------|------------|--------------|--------------|-------------|
+	|unbindable   |shared	   |**unbindable  |    private   | unbindable  |
+	------------------------------------------------------------------------
+
+	* if the shared mount is the only mount in its peer group, making it
+	slave, makes it private automatically. Note that there is no master to
+	which it can be slaved to.
+
+	** slaving a non-shared mount has no effect on the mount.
+
+	Apart from the commands listed below, the 'move' operation also changes
+	the state of a mount depending on type of the destination mount. Its
+	explained in section 5d.
+
+5b) Bind semantics
+
+	Consider the following command
+
+	mount --bind A/a  B/b
+
+	where 'A' is the source mount, 'a' is the dentry in the mount 'A', 'B'
+	is the destination mount and 'b' is the dentry in the destination mount.
+
+	The outcome depends on the type of mount of 'A' and 'B'. The table
+	below contains quick reference.
+   ---------------------------------------------------------------------------
+   |         BIND MOUNT OPERATION                                            |
+   |**************************************************************************
+   |source(A)->| shared       |       private  |       slave    | unbindable |
+   | dest(B)  |               |                |                |            |
+   |   |      |               |                |                |            |
+   |   v      |               |                |                |            |
+   |**************************************************************************
+   |  shared  | shared        |     shared     | shared & slave |  invalid   |
+   |          |               |                |                |            |
+   |non-shared| shared        |      private   |      slave     |  invalid   |
+   ***************************************************************************
+
+     	Details:
+
+	1. 'A' is a shared mount and 'B' is a shared mount. A new mount 'C'
+	which is clone of 'A', is created. Its root dentry is 'a' . 'C' is
+	mounted on mount 'B' at dentry 'b'. Also new mount 'C1', 'C2', 'C3' ...
+	are created and mounted at the dentry 'b' on all mounts where 'B'
+	propagates to. A new propagation tree containing 'C1',..,'Cn' is
+	created. This propagation tree is identical to the propagation tree of
+	'B'.  And finally the peer-group of 'C' is merged with the peer group
+	of 'A'.
+
+	2. 'A' is a private mount and 'B' is a shared mount. A new mount 'C'
+	which is clone of 'A', is created. Its root dentry is 'a'. 'C' is
+	mounted on mount 'B' at dentry 'b'. Also new mount 'C1', 'C2', 'C3' ...
+	are created and mounted at the dentry 'b' on all mounts where 'B'
+	propagates to. A new propagation tree is set containing all new mounts
+	'C', 'C1', .., 'Cn' with exactly the same configuration as the
+	propagation tree for 'B'.
+
+	3. 'A' is a slave mount of mount 'Z' and 'B' is a shared mount. A new
+	mount 'C' which is clone of 'A', is created. Its root dentry is 'a' .
+	'C' is mounted on mount 'B' at dentry 'b'. Also new mounts 'C1', 'C2',
+	'C3' ... are created and mounted at the dentry 'b' on all mounts where
+	'B' propagates to. A new propagation tree containing the new mounts
+	'C','C1',..  'Cn' is created. This propagation tree is identical to the
+	propagation tree for 'B'. And finally the mount 'C' and its peer group
+	is made the slave of mount 'Z'.  In other words, mount 'C' is in the
+	state 'slave and shared'.
+
+	4. 'A' is a unbindable mount and 'B' is a shared mount. This is a
+	invalid operation.
+
+	5. 'A' is a private mount and 'B' is a non-shared(private or slave or
+	unbindable) mount. A new mount 'C' which is clone of 'A', is created.
+	Its root dentry is 'a'. 'C' is mounted on mount 'B' at dentry 'b'.
+
+	6. 'A' is a shared mount and 'B' is a non-shared mount. A new mount 'C'
+	which is a clone of 'A' is created. Its root dentry is 'a'. 'C' is
+	mounted on mount 'B' at dentry 'b'.  'C' is made a member of the
+	peer-group of 'A'.
+
+	7. 'A' is a slave mount of mount 'Z' and 'B' is a non-shared mount. A
+	new mount 'C' which is a clone of 'A' is created. Its root dentry is
+	'a'.  'C' is mounted on mount 'B' at dentry 'b'. Also 'C' is set as a
+	slave mount of 'Z'. In other words 'A' and 'C' are both slave mounts of
+	'Z'.  All mount/unmount events on 'Z' propagates to 'A' and 'C'. But
+	mount/unmount on 'A' do not propagate anywhere else. Similarly
+	mount/unmount on 'C' do not propagate anywhere else.
+
+	8. 'A' is a unbindable mount and 'B' is a non-shared mount. This is a
+	invalid operation. A unbindable mount cannot be bind mounted.
+
+5c) Rbind semantics
+
+	rbind is same as bind. Bind replicates the specified mount.  Rbind
+	replicates all the mounts in the tree belonging to the specified mount.
+	Rbind mount is bind mount applied to all the mounts in the tree.
+
+	If the source tree that is rbind has some unbindable mounts,
+	then the subtree under the unbindable mount is pruned in the new
+	location.
+
+	eg: let's say we have the following mount tree.
+
+		A
+	      /   \
+	      B   C
+	     / \ / \
+	     D E F G
+
+	     Let's say all the mount except the mount C in the tree are
+	     of a type other than unbindable.
+
+	     If this tree is rbound to say Z
+
+	     We will have the following tree at the new location.
+
+		Z
+		|
+		A'
+	       /
+	      B'		Note how the tree under C is pruned
+	     / \ 		in the new location.
+	    D' E'
+
+
+
+5d) Move semantics
+
+	Consider the following command
+
+	mount --move A  B/b
+
+	where 'A' is the source mount, 'B' is the destination mount and 'b' is
+	the dentry in the destination mount.
+
+	The outcome depends on the type of the mount of 'A' and 'B'. The table
+	below is a quick reference.
+   ---------------------------------------------------------------------------
+   |         		MOVE MOUNT OPERATION                                 |
+   |**************************************************************************
+   | source(A)->| shared      |       private  |       slave    | unbindable |
+   | dest(B)  |               |                |                |            |
+   |   |      |               |                |                |            |
+   |   v      |               |                |                |            |
+   |**************************************************************************
+   |  shared  | shared        |     shared     |shared and slave|  invalid   |
+   |          |               |                |                |            |
+   |non-shared| shared        |      private   |    slave       | unbindable |
+   ***************************************************************************
+	NOTE: moving a mount residing under a shared mount is invalid.
+
+      Details follow:
+
+	1. 'A' is a shared mount and 'B' is a shared mount.  The mount 'A' is
+	mounted on mount 'B' at dentry 'b'.  Also new mounts 'A1', 'A2'...'An'
+	are created and mounted at dentry 'b' on all mounts that receive
+	propagation from mount 'B'. A new propagation tree is created in the
+	exact same configuration as that of 'B'. This new propagation tree
+	contains all the new mounts 'A1', 'A2'...  'An'.  And this new
+	propagation tree is appended to the already existing propagation tree
+	of 'A'.
+
+	2. 'A' is a private mount and 'B' is a shared mount. The mount 'A' is
+	mounted on mount 'B' at dentry 'b'. Also new mount 'A1', 'A2'... 'An'
+	are created and mounted at dentry 'b' on all mounts that receive
+	propagation from mount 'B'. The mount 'A' becomes a shared mount and a
+	propagation tree is created which is identical to that of
+	'B'. This new propagation tree contains all the new mounts 'A1',
+	'A2'...  'An'.
+
+	3. 'A' is a slave mount of mount 'Z' and 'B' is a shared mount.  The
+	mount 'A' is mounted on mount 'B' at dentry 'b'.  Also new mounts 'A1',
+	'A2'... 'An' are created and mounted at dentry 'b' on all mounts that
+	receive propagation from mount 'B'. A new propagation tree is created
+	in the exact same configuration as that of 'B'. This new propagation
+	tree contains all the new mounts 'A1', 'A2'...  'An'.  And this new
+	propagation tree is appended to the already existing propagation tree of
+	'A'.  Mount 'A' continues to be the slave mount of 'Z' but it also
+	becomes 'shared'.
+
+	4. 'A' is a unbindable mount and 'B' is a shared mount. The operation
+	is invalid. Because mounting anything on the shared mount 'B' can
+	create new mounts that get mounted on the mounts that receive
+	propagation from 'B'.  And since the mount 'A' is unbindable, cloning
+	it to mount at other mountpoints is not possible.
+
+	5. 'A' is a private mount and 'B' is a non-shared(private or slave or
+	unbindable) mount. The mount 'A' is mounted on mount 'B' at dentry 'b'.
+
+	6. 'A' is a shared mount and 'B' is a non-shared mount.  The mount 'A'
+	is mounted on mount 'B' at dentry 'b'.  Mount 'A' continues to be a
+	shared mount.
+
+	7. 'A' is a slave mount of mount 'Z' and 'B' is a non-shared mount.
+	The mount 'A' is mounted on mount 'B' at dentry 'b'.  Mount 'A'
+	continues to be a slave mount of mount 'Z'.
+
+	8. 'A' is a unbindable mount and 'B' is a non-shared mount. The mount
+	'A' is mounted on mount 'B' at dentry 'b'. Mount 'A' continues to be a
+	unbindable mount.
+
+5e) Mount semantics
+
+	Consider the following command
+
+	mount device  B/b
+
+	'B' is the destination mount and 'b' is the dentry in the destination
+	mount.
+
+	The above operation is the same as bind operation with the exception
+	that the source mount is always a private mount.
+
+
+5f) Unmount semantics
+
+	Consider the following command
+
+	umount A
+
+	where 'A' is a mount mounted on mount 'B' at dentry 'b'.
+
+	If mount 'B' is shared, then all most-recently-mounted mounts at dentry
+	'b' on mounts that receive propagation from mount 'B' and does not have
+	sub-mounts within them are unmounted.
+
+	Example: Let's say 'B1', 'B2', 'B3' are shared mounts that propagate to
+	each other.
+
+	let's say 'A1', 'A2', 'A3' are first mounted at dentry 'b' on mount
+	'B1', 'B2' and 'B3' respectively.
+
+	let's say 'C1', 'C2', 'C3' are next mounted at the same dentry 'b' on
+	mount 'B1', 'B2' and 'B3' respectively.
+
+	if 'C1' is unmounted, all the mounts that are most-recently-mounted on
+	'B1' and on the mounts that 'B1' propagates-to are unmounted.
+
+	'B1' propagates to 'B2' and 'B3'. And the most recently mounted mount
+	on 'B2' at dentry 'b' is 'C2', and that of mount 'B3' is 'C3'.
+
+	So all 'C1', 'C2' and 'C3' should be unmounted.
+
+	If any of 'C2' or 'C3' has some child mounts, then that mount is not
+	unmounted, but all other mounts are unmounted. However if 'C1' is told
+	to be unmounted and 'C1' has some sub-mounts, the umount operation is
+	failed entirely.
+
+5g) Clone Namespace
+
+	A cloned namespace contains all the mounts as that of the parent
+	namespace.
+
+	Let's say 'A' and 'B' are the corresponding mounts in the parent and the
+	child namespace.
+
+	If 'A' is shared, then 'B' is also shared and 'A' and 'B' propagate to
+	each other.
+
+	If 'A' is a slave mount of 'Z', then 'B' is also the slave mount of
+	'Z'.
+
+	If 'A' is a private mount, then 'B' is a private mount too.
+
+	If 'A' is unbindable mount, then 'B' is a unbindable mount too.
+
+
+6) Quiz
+
+	A. What is the result of the following command sequence?
+
+		mount --bind /mnt /mnt
+		mount --make-shared /mnt
+		mount --bind /mnt /tmp
+		mount --move /tmp /mnt/1
+
+		what should be the contents of /mnt /mnt/1 /mnt/1/1 should be?
+		Should they all be identical? or should /mnt and /mnt/1 be
+		identical only?
+
+
+	B. What is the result of the following command sequence?
+
+		mount --make-rshared /
+		mkdir -p /v/1
+		mount --rbind / /v/1
+
+		what should be the content of /v/1/v/1 be?
+
+
+	C. What is the result of the following command sequence?
+
+		mount --bind /mnt /mnt
+		mount --make-shared /mnt
+		mkdir -p /mnt/1/2/3 /mnt/1/test
+		mount --bind /mnt/1 /tmp
+		mount --make-slave /mnt
+		mount --make-shared /mnt
+		mount --bind /mnt/1/2 /tmp1
+		mount --make-slave /mnt
+
+		At this point we have the first mount at /tmp and
+		its root dentry is 1. Let's call this mount 'A'
+		And then we have a second mount at /tmp1 with root
+		dentry 2. Let's call this mount 'B'
+		Next we have a third mount at /mnt with root dentry
+		mnt. Let's call this mount 'C'
+
+		'B' is the slave of 'A' and 'C' is a slave of 'B'
+		A -> B -> C
+
+		at this point if we execute the following command
+
+		mount --bind /bin /tmp/test
+
+		The mount is attempted on 'A'
+
+		will the mount propagate to 'B' and 'C' ?
+
+		what would be the contents of
+		/mnt/1/test be?
+
+7) FAQ
+
+	Q1. Why is bind mount needed? How is it different from symbolic links?
+		symbolic links can get stale if the destination mount gets
+		unmounted or moved. Bind mounts continue to exist even if the
+		other mount is unmounted or moved.
+
+	Q2. Why can't the shared subtree be implemented using exportfs?
+
+		exportfs is a heavyweight way of accomplishing part of what
+		shared subtree can do. I cannot imagine a way to implement the
+		semantics of slave mount using exportfs?
+
+	Q3 Why is unbindable mount needed?
+
+		Let's say we want to replicate the mount tree at multiple
+		locations within the same subtree.
+
+		if one rbind mounts a tree within the same subtree 'n' times
+		the number of mounts created is an exponential function of 'n'.
+		Having unbindable mount can help prune the unneeded bind
+		mounts. Here is an example.
+
+		step 1:
+		   let's say the root tree has just two directories with
+		   one vfsmount.
+				    root
+				   /    \
+				  tmp    usr
+
+		    And we want to replicate the tree at multiple
+		    mountpoints under /root/tmp
+
+		step2:
+		      mount --make-shared /root
+
+		      mkdir -p /tmp/m1
+
+		      mount --rbind /root /tmp/m1
+
+		      the new tree now looks like this:
+
+				    root
+				   /    \
+				 tmp    usr
+				/
+			       m1
+			      /  \
+			     tmp  usr
+			     /
+			    m1
+
+			  it has two vfsmounts
+
+		step3:
+			    mkdir -p /tmp/m2
+			    mount --rbind /root /tmp/m2
+
+			the new tree now looks like this:
+
+				      root
+				     /    \
+				   tmp     usr
+				  /    \
+				m1       m2
+			       / \       /  \
+			     tmp  usr   tmp  usr
+			     / \          /
+			    m1  m2      m1
+				/ \     /  \
+			      tmp usr  tmp   usr
+			      /        / \
+			     m1       m1  m2
+			    /  \
+			  tmp   usr
+			  /  \
+			 m1   m2
+
+		       it has 6 vfsmounts
+
+		step 4:
+			  mkdir -p /tmp/m3
+			  mount --rbind /root /tmp/m3
+
+			  I won't draw the tree..but it has 24 vfsmounts
+
+
+		at step i the number of vfsmounts is V[i] = i*V[i-1].
+		This is an exponential function. And this tree has way more
+		mounts than what we really needed in the first place.
+
+		One could use a series of umount at each step to prune
+		out the unneeded mounts. But there is a better solution.
+		Unclonable mounts come in handy here.
+
+		step 1:
+		   let's say the root tree has just two directories with
+		   one vfsmount.
+				    root
+				   /    \
+				  tmp    usr
+
+		    How do we set up the same tree at multiple locations under
+		    /root/tmp
+
+		step2:
+		      mount --bind /root/tmp /root/tmp
+
+		      mount --make-rshared /root
+		      mount --make-unbindable /root/tmp
+
+		      mkdir -p /tmp/m1
+
+		      mount --rbind /root /tmp/m1
+
+		      the new tree now looks like this:
+
+				    root
+				   /    \
+				 tmp    usr
+				/
+			       m1
+			      /  \
+			     tmp  usr
+
+		step3:
+			    mkdir -p /tmp/m2
+			    mount --rbind /root /tmp/m2
+
+		      the new tree now looks like this:
+
+				    root
+				   /    \
+				 tmp    usr
+				/   \
+			       m1     m2
+			      /  \     / \
+			     tmp  usr tmp usr
+
+		step4:
+
+			    mkdir -p /tmp/m3
+			    mount --rbind /root /tmp/m3
+
+		      the new tree now looks like this:
+
+				    	  root
+				      /    	  \
+				     tmp    	   usr
+			         /    \    \
+			       m1     m2     m3
+			      /  \     / \    /  \
+			     tmp  usr tmp usr tmp usr
+
+8) Implementation
+
+8A) Datastructure
+
+	4 new fields are introduced to struct vfsmount
+	->mnt_share
+	->mnt_slave_list
+	->mnt_slave
+	->mnt_master
+
+	->mnt_share links together all the mount to/from which this vfsmount
+		send/receives propagation events.
+
+	->mnt_slave_list links all the mounts to which this vfsmount propagates
+		to.
+
+	->mnt_slave links together all the slaves that its master vfsmount
+		propagates to.
+
+	->mnt_master points to the master vfsmount from which this vfsmount
+		receives propagation.
+
+	->mnt_flags takes two more flags to indicate the propagation status of
+		the vfsmount.  MNT_SHARE indicates that the vfsmount is a shared
+		vfsmount.  MNT_UNCLONABLE indicates that the vfsmount cannot be
+		replicated.
+
+	All the shared vfsmounts in a peer group form a cyclic list through
+	->mnt_share.
+
+	All vfsmounts with the same ->mnt_master form on a cyclic list anchored
+	in ->mnt_master->mnt_slave_list and going through ->mnt_slave.
+
+	 ->mnt_master can point to arbitrary (and possibly different) members
+	 of master peer group.  To find all immediate slaves of a peer group
+	 you need to go through _all_ ->mnt_slave_list of its members.
+	 Conceptually it's just a single set - distribution among the
+	 individual lists does not affect propagation or the way propagation
+	 tree is modified by operations.
+
+	All vfsmounts in a peer group have the same ->mnt_master.  If it is
+	non-NULL, they form a contiguous (ordered) segment of slave list.
+
+	A example propagation tree looks as shown in the figure below.
+	[ NOTE: Though it looks like a forest, if we consider all the shared
+	mounts as a conceptual entity called 'pnode', it becomes a tree]
+
+
+		        A <--> B <--> C <---> D
+		       /|\	      /|      |\
+		      / F G	     J K      H I
+		     /
+		    E<-->K
+			/|\
+		       M L N
+
+	In the above figure  A,B,C and D all are shared and propagate to each
+	other.   'A' has got 3 slave mounts 'E' 'F' and 'G' 'C' has got 2 slave
+	mounts 'J' and 'K'  and  'D' has got two slave mounts 'H' and 'I'.
+	'E' is also shared with 'K' and they propagate to each other.  And
+	'K' has 3 slaves 'M', 'L' and 'N'
+
+	A's ->mnt_share links with the ->mnt_share of 'B' 'C' and 'D'
+
+	A's ->mnt_slave_list links with ->mnt_slave of 'E', 'K', 'F' and 'G'
+
+	E's ->mnt_share links with ->mnt_share of K
+	'E', 'K', 'F', 'G' have their ->mnt_master point to struct
+				vfsmount of 'A'
+	'M', 'L', 'N' have their ->mnt_master point to struct vfsmount of 'K'
+	K's ->mnt_slave_list links with ->mnt_slave of 'M', 'L' and 'N'
+
+	C's ->mnt_slave_list links with ->mnt_slave of 'J' and 'K'
+	J and K's ->mnt_master points to struct vfsmount of C
+	and finally D's ->mnt_slave_list links with ->mnt_slave of 'H' and 'I'
+	'H' and 'I' have their ->mnt_master pointing to struct vfsmount of 'D'.
+
+
+	NOTE: The propagation tree is orthogonal to the mount tree.
+
+8B Locking:
+
+	->mnt_share, ->mnt_slave, ->mnt_slave_list, ->mnt_master are protected
+	by namespace_sem (exclusive for modifications, shared for reading).
+
+	Normally we have ->mnt_flags modifications serialized by vfsmount_lock.
+	There are two exceptions: do_add_mount() and clone_mnt().
+	The former modifies a vfsmount that has not been visible in any shared
+	data structures yet.
+	The latter holds namespace_sem and the only references to vfsmount
+	are in lists that can't be traversed without namespace_sem.
+
+8C Algorithm:
+
+	The crux of the implementation resides in rbind/move operation.
+
+	The overall algorithm breaks the operation into 3 phases: (look at
+	attach_recursive_mnt() and propagate_mnt())
+
+	1. prepare phase.
+	2. commit phases.
+	3. abort phases.
+
+	Prepare phase:
+
+	for each mount in the source tree:
+		   a) Create the necessary number of mount trees to
+		   	be attached to each of the mounts that receive
+			propagation from the destination mount.
+		   b) Do not attach any of the trees to its destination.
+		      However note down its ->mnt_parent and ->mnt_mountpoint
+		   c) Link all the new mounts to form a propagation tree that
+		      is identical to the propagation tree of the destination
+		      mount.
+
+		   If this phase is successful, there should be 'n' new
+		   propagation trees; where 'n' is the number of mounts in the
+		   source tree.  Go to the commit phase
+
+		   Also there should be 'm' new mount trees, where 'm' is
+		   the number of mounts to which the destination mount
+		   propagates to.
+
+		   if any memory allocations fail, go to the abort phase.
+
+	Commit phase
+		attach each of the mount trees to their corresponding
+		destination mounts.
+
+	Abort phase
+		delete all the newly created trees.
+
+	NOTE: all the propagation related functionality resides in the file
+	pnode.c
+
+
+------------------------------------------------------------------------
+
+version 0.1  (created the initial document, Ram Pai linuxram@us.ibm.com)
+version 0.2  (Incorporated comments from Al Viro)
diff --git a/Documentation/filesystems/spufs.txt b/Documentation/filesystems/spufs.txt
new file mode 100644
index 000000000..1343d118a
--- /dev/null
+++ b/Documentation/filesystems/spufs.txt
@@ -0,0 +1,521 @@
+SPUFS(2)                   Linux Programmer's Manual                  SPUFS(2)
+
+
+
+NAME
+       spufs - the SPU file system
+
+
+DESCRIPTION
+       The SPU file system is used on PowerPC machines that implement the Cell
+       Broadband Engine Architecture in order to access Synergistic  Processor
+       Units (SPUs).
+
+       The file system provides a name space similar to posix shared memory or
+       message queues. Users that have write permissions on  the  file  system
+       can use spu_create(2) to establish SPU contexts in the spufs root.
+
+       Every SPU context is represented by a directory containing a predefined
+       set of files. These files can be used for manipulating the state of the
+       logical SPU. Users can change permissions on those files, but not actu-
+       ally add or remove files.
+
+
+MOUNT OPTIONS
+       uid=<uid>
+              set the user owning the mount point, the default is 0 (root).
+
+       gid=<gid>
+              set the group owning the mount point, the default is 0 (root).
+
+
+FILES
+       The files in spufs mostly follow the standard behavior for regular sys-
+       tem  calls like read(2) or write(2), but often support only a subset of
+       the operations supported on regular file systems. This list details the
+       supported  operations  and  the  deviations  from  the behaviour in the
+       respective man pages.
+
+       All files that support the read(2) operation also support readv(2)  and
+       all  files  that support the write(2) operation also support writev(2).
+       All files support the access(2) and stat(2) family of  operations,  but
+       only  the  st_mode,  st_nlink,  st_uid and st_gid fields of struct stat
+       contain reliable information.
+
+       All files support the chmod(2)/fchmod(2) and chown(2)/fchown(2)  opera-
+       tions,  but  will  not be able to grant permissions that contradict the
+       possible operations, e.g. read access on the wbox file.
+
+       The current set of files is:
+
+
+   /mem
+       the contents of the local storage memory  of  the  SPU.   This  can  be
+       accessed  like  a regular shared memory file and contains both code and
+       data in the address space of the SPU.  The possible  operations  on  an
+       open mem file are:
+
+       read(2), pread(2), write(2), pwrite(2), lseek(2)
+              These  operate  as  documented, with the exception that seek(2),
+              write(2) and pwrite(2) are not supported beyond the end  of  the
+              file. The file size is the size of the local storage of the SPU,
+              which normally is 256 kilobytes.
+
+       mmap(2)
+              Mapping mem into the process address space gives access  to  the
+              SPU  local  storage  within  the  process  address  space.  Only
+              MAP_SHARED mappings are allowed.
+
+
+   /mbox
+       The first SPU to CPU communication mailbox. This file is read-only  and
+       can  be  read  in  units of 32 bits.  The file can only be used in non-
+       blocking mode and it even poll() will not block on  it.   The  possible
+       operations on an open mbox file are:
+
+       read(2)
+              If  a  count smaller than four is requested, read returns -1 and
+              sets errno to EINVAL.  If there is no data available in the mail
+              box,  the  return  value  is set to -1 and errno becomes EAGAIN.
+              When data has been read successfully, four bytes are  placed  in
+              the data buffer and the value four is returned.
+
+
+   /ibox
+       The  second  SPU  to CPU communication mailbox. This file is similar to
+       the first mailbox file, but can be read in blocking I/O mode,  and  the
+       poll  family of system calls can be used to wait for it.  The  possible
+       operations on an open ibox file are:
+
+       read(2)
+              If a count smaller than four is requested, read returns  -1  and
+              sets errno to EINVAL.  If there is no data available in the mail
+              box and the file descriptor has been opened with O_NONBLOCK, the
+              return value is set to -1 and errno becomes EAGAIN.
+
+              If  there  is  no  data  available  in the mail box and the file
+              descriptor has been opened without  O_NONBLOCK,  the  call  will
+              block  until  the  SPU  writes to its interrupt mailbox channel.
+              When data has been read successfully, four bytes are  placed  in
+              the data buffer and the value four is returned.
+
+       poll(2)
+              Poll  on  the  ibox  file returns (POLLIN | POLLRDNORM) whenever
+              data is available for reading.
+
+
+   /wbox
+       The CPU to SPU communation mailbox. It is write-only and can be written
+       in  units  of  32  bits. If the mailbox is full, write() will block and
+       poll can be used to wait for it becoming  empty  again.   The  possible
+       operations  on  an open wbox file are: write(2) If a count smaller than
+       four is requested, write returns -1 and sets errno to EINVAL.  If there
+       is  no space available in the mail box and the file descriptor has been
+       opened with O_NONBLOCK, the return value is set to -1 and errno becomes
+       EAGAIN.
+
+       If  there is no space available in the mail box and the file descriptor
+       has been opened without O_NONBLOCK, the call will block until  the  SPU
+       reads  from  its PPE mailbox channel.  When data has been read success-
+       fully, four bytes are placed in the data buffer and the value  four  is
+       returned.
+
+       poll(2)
+              Poll  on  the  ibox file returns (POLLOUT | POLLWRNORM) whenever
+              space is available for writing.
+
+
+   /mbox_stat
+   /ibox_stat
+   /wbox_stat
+       Read-only files that contain the length of the current queue, i.e.  how
+       many  words  can  be  read  from  mbox or ibox or how many words can be
+       written to wbox without blocking.  The files can be read only in 4-byte
+       units  and  return  a  big-endian  binary integer number.  The possible
+       operations on an open *box_stat file are:
+
+       read(2)
+              If a count smaller than four is requested, read returns  -1  and
+              sets errno to EINVAL.  Otherwise, a four byte value is placed in
+              the data buffer, containing the number of elements that  can  be
+              read  from  (for  mbox_stat  and  ibox_stat)  or written to (for
+              wbox_stat) the respective mail box without blocking or resulting
+              in EAGAIN.
+
+
+   /npc
+   /decr
+   /decr_status
+   /spu_tag_mask
+   /event_mask
+   /srr0
+       Internal  registers  of  the SPU. The representation is an ASCII string
+       with the numeric value of the next instruction to  be  executed.  These
+       can  be  used in read/write mode for debugging, but normal operation of
+       programs should not rely on them because access to any of  them  except
+       npc requires an SPU context save and is therefore very inefficient.
+
+       The contents of these files are:
+
+       npc                 Next Program Counter
+
+       decr                SPU Decrementer
+
+       decr_status         Decrementer Status
+
+       spu_tag_mask        MFC tag mask for SPU DMA
+
+       event_mask          Event mask for SPU interrupts
+
+       srr0                Interrupt Return address register
+
+
+       The   possible   operations   on   an   open  npc,  decr,  decr_status,
+       spu_tag_mask, event_mask or srr0 file are:
+
+       read(2)
+              When the count supplied to the read call  is  shorter  than  the
+              required  length for the pointer value plus a newline character,
+              subsequent reads from the same file descriptor  will  result  in
+              completing  the string, regardless of changes to the register by
+              a running SPU task.  When a complete string has been  read,  all
+              subsequent read operations will return zero bytes and a new file
+              descriptor needs to be opened to read the value again.
+
+       write(2)
+              A write operation on the file results in setting the register to
+              the  value  given  in  the string. The string is parsed from the
+              beginning to the first non-numeric character or the end  of  the
+              buffer.  Subsequent writes to the same file descriptor overwrite
+              the previous setting.
+
+
+   /fpcr
+       This file gives access to the Floating Point Status and Control  Regis-
+       ter as a four byte long file. The operations on the fpcr file are:
+
+       read(2)
+              If  a  count smaller than four is requested, read returns -1 and
+              sets errno to EINVAL.  Otherwise, a four byte value is placed in
+              the data buffer, containing the current value of the fpcr regis-
+              ter.
+
+       write(2)
+              If a count smaller than four is requested, write returns -1  and
+              sets  errno  to  EINVAL.  Otherwise, a four byte value is copied
+              from the data buffer, updating the value of the fpcr register.
+
+
+   /signal1
+   /signal2
+       The two signal notification channels of an SPU.  These  are  read-write
+       files  that  operate  on  a 32 bit word.  Writing to one of these files
+       triggers an interrupt on the SPU.  The  value  written  to  the  signal
+       files can be read from the SPU through a channel read or from host user
+       space through the file.  After the value has been read by the  SPU,  it
+       is  reset  to zero.  The possible operations on an open signal1 or sig-
+       nal2 file are:
+
+       read(2)
+              If a count smaller than four is requested, read returns  -1  and
+              sets errno to EINVAL.  Otherwise, a four byte value is placed in
+              the data buffer, containing the current value of  the  specified
+              signal notification register.
+
+       write(2)
+              If  a count smaller than four is requested, write returns -1 and
+              sets errno to EINVAL.  Otherwise, a four byte  value  is  copied
+              from the data buffer, updating the value of the specified signal
+              notification register.  The signal  notification  register  will
+              either be replaced with the input data or will be updated to the
+              bitwise OR or the old value and the input data, depending on the
+              contents  of  the  signal1_type,  or  signal2_type respectively,
+              file.
+
+
+   /signal1_type
+   /signal2_type
+       These two files change the behavior of the signal1 and signal2  notifi-
+       cation  files.  The  contain  a numerical ASCII string which is read as
+       either "1" or "0".  In mode 0 (overwrite), the  hardware  replaces  the
+       contents of the signal channel with the data that is written to it.  in
+       mode 1 (logical OR), the hardware accumulates the bits that are  subse-
+       quently written to it.  The possible operations on an open signal1_type
+       or signal2_type file are:
+
+       read(2)
+              When the count supplied to the read call  is  shorter  than  the
+              required  length  for the digit plus a newline character, subse-
+              quent reads from the same file descriptor will  result  in  com-
+              pleting  the  string.  When a complete string has been read, all
+              subsequent read operations will return zero bytes and a new file
+              descriptor needs to be opened to read the value again.
+
+       write(2)
+              A write operation on the file results in setting the register to
+              the value given in the string. The string  is  parsed  from  the
+              beginning  to  the first non-numeric character or the end of the
+              buffer.  Subsequent writes to the same file descriptor overwrite
+              the previous setting.
+
+
+EXAMPLES
+       /etc/fstab entry
+              none      /spu      spufs     gid=spu   0    0
+
+
+AUTHORS
+       Arnd  Bergmann  <arndb@de.ibm.com>,  Mark  Nutter <mnutter@us.ibm.com>,
+       Ulrich Weigand <Ulrich.Weigand@de.ibm.com>
+
+SEE ALSO
+       capabilities(7), close(2), spu_create(2), spu_run(2), spufs(7)
+
+
+
+Linux                             2005-09-28                          SPUFS(2)
+
+------------------------------------------------------------------------------
+
+SPU_RUN(2)                 Linux Programmer's Manual                SPU_RUN(2)
+
+
+
+NAME
+       spu_run - execute an spu context
+
+
+SYNOPSIS
+       #include <sys/spu.h>
+
+       int spu_run(int fd, unsigned int *npc, unsigned int *event);
+
+DESCRIPTION
+       The  spu_run system call is used on PowerPC machines that implement the
+       Cell Broadband Engine Architecture in order to access Synergistic  Pro-
+       cessor  Units  (SPUs).  It  uses the fd that was returned from spu_cre-
+       ate(2) to address a specific SPU context. When the context gets  sched-
+       uled  to a physical SPU, it starts execution at the instruction pointer
+       passed in npc.
+
+       Execution of SPU code happens synchronously, meaning that spu_run  does
+       not  return  while the SPU is still running. If there is a need to exe-
+       cute SPU code in parallel with other code on either  the  main  CPU  or
+       other  SPUs,  you  need to create a new thread of execution first, e.g.
+       using the pthread_create(3) call.
+
+       When spu_run returns, the current value of the SPU instruction  pointer
+       is  written back to npc, so you can call spu_run again without updating
+       the pointers.
+
+       event can be a NULL pointer or point to an extended  status  code  that
+       gets  filled  when spu_run returns. It can be one of the following con-
+       stants:
+
+       SPE_EVENT_DMA_ALIGNMENT
+              A DMA alignment error
+
+       SPE_EVENT_SPE_DATA_SEGMENT
+              A DMA segmentation error
+
+       SPE_EVENT_SPE_DATA_STORAGE
+              A DMA storage error
+
+       If NULL is passed as the event argument, these errors will result in  a
+       signal delivered to the calling process.
+
+RETURN VALUE
+       spu_run  returns the value of the spu_status register or -1 to indicate
+       an error and set errno to one of the error  codes  listed  below.   The
+       spu_status  register  value  contains  a  bit  mask of status codes and
+       optionally a 14 bit code returned from the stop-and-signal  instruction
+       on the SPU. The bit masks for the status codes are:
+
+       0x02   SPU was stopped by stop-and-signal.
+
+       0x04   SPU was stopped by halt.
+
+       0x08   SPU is waiting for a channel.
+
+       0x10   SPU is in single-step mode.
+
+       0x20   SPU has tried to execute an invalid instruction.
+
+       0x40   SPU has tried to access an invalid channel.
+
+       0x3fff0000
+              The  bits  masked with this value contain the code returned from
+              stop-and-signal.
+
+       There are always one or more of the lower eight bits set  or  an  error
+       code is returned from spu_run.
+
+ERRORS
+       EAGAIN or EWOULDBLOCK
+              fd is in non-blocking mode and spu_run would block.
+
+       EBADF  fd is not a valid file descriptor.
+
+       EFAULT npc is not a valid pointer or status is neither NULL nor a valid
+              pointer.
+
+       EINTR  A signal occurred while spu_run was in progress.  The npc  value
+              has  been updated to the new program counter value if necessary.
+
+       EINVAL fd is not a file descriptor returned from spu_create(2).
+
+       ENOMEM Insufficient memory was available to handle a page fault result-
+              ing from an MFC direct memory access.
+
+       ENOSYS the functionality is not provided by the current system, because
+              either the hardware does not provide SPUs or the spufs module is
+              not loaded.
+
+
+NOTES
+       spu_run  is  meant  to  be  used  from  libraries that implement a more
+       abstract interface to SPUs, not to be used from  regular  applications.
+       See  http://www.bsc.es/projects/deepcomputing/linuxoncell/ for the rec-
+       ommended libraries.
+
+
+CONFORMING TO
+       This call is Linux specific and only implemented by the ppc64 architec-
+       ture. Programs using this system call are not portable.
+
+
+BUGS
+       The code does not yet fully implement all features lined out here.
+
+
+AUTHOR
+       Arnd Bergmann <arndb@de.ibm.com>
+
+SEE ALSO
+       capabilities(7), close(2), spu_create(2), spufs(7)
+
+
+
+Linux                             2005-09-28                        SPU_RUN(2)
+
+------------------------------------------------------------------------------
+
+SPU_CREATE(2)              Linux Programmer's Manual             SPU_CREATE(2)
+
+
+
+NAME
+       spu_create - create a new spu context
+
+
+SYNOPSIS
+       #include <sys/types.h>
+       #include <sys/spu.h>
+
+       int spu_create(const char *pathname, int flags, mode_t mode);
+
+DESCRIPTION
+       The  spu_create  system call is used on PowerPC machines that implement
+       the Cell Broadband Engine Architecture in order to  access  Synergistic
+       Processor  Units (SPUs). It creates a new logical context for an SPU in
+       pathname and returns a handle to associated  with  it.   pathname  must
+       point  to  a  non-existing directory in the mount point of the SPU file
+       system (spufs).  When spu_create is successful, a directory  gets  cre-
+       ated on pathname and it is populated with files.
+
+       The  returned  file  handle can only be passed to spu_run(2) or closed,
+       other operations are not defined on it. When it is closed, all  associ-
+       ated  directory entries in spufs are removed. When the last file handle
+       pointing either inside  of  the  context  directory  or  to  this  file
+       descriptor is closed, the logical SPU context is destroyed.
+
+       The  parameter flags can be zero or any bitwise or'd combination of the
+       following constants:
+
+       SPU_RAWIO
+              Allow mapping of some of the hardware registers of the SPU  into
+              user space. This flag requires the CAP_SYS_RAWIO capability, see
+              capabilities(7).
+
+       The mode parameter specifies the permissions used for creating the  new
+       directory  in  spufs.   mode is modified with the user's umask(2) value
+       and then used for both the directory and the files contained in it. The
+       file permissions mask out some more bits of mode because they typically
+       support only read or write access. See stat(2) for a full list  of  the
+       possible mode values.
+
+
+RETURN VALUE
+       spu_create  returns a new file descriptor. It may return -1 to indicate
+       an error condition and set errno to  one  of  the  error  codes  listed
+       below.
+
+
+ERRORS
+       EACCESS
+              The  current  user does not have write access on the spufs mount
+              point.
+
+       EEXIST An SPU context already exists at the given path name.
+
+       EFAULT pathname is not a valid string pointer in  the  current  address
+              space.
+
+       EINVAL pathname is not a directory in the spufs mount point.
+
+       ELOOP  Too many symlinks were found while resolving pathname.
+
+       EMFILE The process has reached its maximum open file limit.
+
+       ENAMETOOLONG
+              pathname was too long.
+
+       ENFILE The system has reached the global open file limit.
+
+       ENOENT Part of pathname could not be resolved.
+
+       ENOMEM The kernel could not allocate all resources required.
+
+       ENOSPC There  are  not  enough  SPU resources available to create a new
+              context or the user specific limit for the number  of  SPU  con-
+              texts has been reached.
+
+       ENOSYS the functionality is not provided by the current system, because
+              either the hardware does not provide SPUs or the spufs module is
+              not loaded.
+
+       ENOTDIR
+              A part of pathname is not a directory.
+
+
+
+NOTES
+       spu_create  is  meant  to  be used from libraries that implement a more
+       abstract interface to SPUs, not to be used from  regular  applications.
+       See  http://www.bsc.es/projects/deepcomputing/linuxoncell/ for the rec-
+       ommended libraries.
+
+
+FILES
+       pathname must point to a location beneath the mount point of spufs.  By
+       convention, it gets mounted in /spu.
+
+
+CONFORMING TO
+       This call is Linux specific and only implemented by the ppc64 architec-
+       ture. Programs using this system call are not portable.
+
+
+BUGS
+       The code does not yet fully implement all features lined out here.
+
+
+AUTHOR
+       Arnd Bergmann <arndb@de.ibm.com>
+
+SEE ALSO
+       capabilities(7), close(2), spu_run(2), spufs(7)
+
+
+
+Linux                             2005-09-28                     SPU_CREATE(2)
diff --git a/Documentation/filesystems/squashfs.txt b/Documentation/filesystems/squashfs.txt
new file mode 100644
index 000000000..e5274f84d
--- /dev/null
+++ b/Documentation/filesystems/squashfs.txt
@@ -0,0 +1,259 @@
+SQUASHFS 4.0 FILESYSTEM
+=======================
+
+Squashfs is a compressed read-only filesystem for Linux.
+It uses zlib, lz4, lzo, or xz compression to compress files, inodes and
+directories.  Inodes in the system are very small and all blocks are packed to
+minimise data overhead. Block sizes greater than 4K are supported up to a
+maximum of 1Mbytes (default block size 128K).
+
+Squashfs is intended for general read-only filesystem use, for archival
+use (i.e. in cases where a .tar.gz file may be used), and in constrained
+block device/memory systems (e.g. embedded systems) where low overhead is
+needed.
+
+Mailing list: squashfs-devel@lists.sourceforge.net
+Web site: www.squashfs.org
+
+1. FILESYSTEM FEATURES
+----------------------
+
+Squashfs filesystem features versus Cramfs:
+
+				Squashfs		Cramfs
+
+Max filesystem size:		2^64			256 MiB
+Max file size:			~ 2 TiB			16 MiB
+Max files:			unlimited		unlimited
+Max directories:		unlimited		unlimited
+Max entries per directory:	unlimited		unlimited
+Max block size:			1 MiB			4 KiB
+Metadata compression:		yes			no
+Directory indexes:		yes			no
+Sparse file support:		yes			no
+Tail-end packing (fragments):	yes			no
+Exportable (NFS etc.):		yes			no
+Hard link support:		yes			no
+"." and ".." in readdir:	yes			no
+Real inode numbers:		yes			no
+32-bit uids/gids:		yes			no
+File creation time:		yes			no
+Xattr support:			yes			no
+ACL support:			no			no
+
+Squashfs compresses data, inodes and directories.  In addition, inode and
+directory data are highly compacted, and packed on byte boundaries.  Each
+compressed inode is on average 8 bytes in length (the exact length varies on
+file type, i.e. regular file, directory, symbolic link, and block/char device
+inodes have different sizes).
+
+2. USING SQUASHFS
+-----------------
+
+As squashfs is a read-only filesystem, the mksquashfs program must be used to
+create populated squashfs filesystems.  This and other squashfs utilities
+can be obtained from http://www.squashfs.org.  Usage instructions can be
+obtained from this site also.
+
+The squashfs-tools development tree is now located on kernel.org
+	git://git.kernel.org/pub/scm/fs/squashfs/squashfs-tools.git
+
+3. SQUASHFS FILESYSTEM DESIGN
+-----------------------------
+
+A squashfs filesystem consists of a maximum of nine parts, packed together on a
+byte alignment:
+
+	 ---------------
+	|  superblock 	|
+	|---------------|
+	|  compression  |
+	|    options    |
+	|---------------|
+	|  datablocks   |
+	|  & fragments  |
+	|---------------|
+	|  inode table	|
+	|---------------|
+	|   directory	|
+	|     table     |
+	|---------------|
+	|   fragment	|
+	|    table      |
+	|---------------|
+	|    export     |
+	|    table      |
+	|---------------|
+	|    uid/gid	|
+	|  lookup table	|
+	|---------------|
+	|     xattr     |
+	|     table	|
+	 ---------------
+
+Compressed data blocks are written to the filesystem as files are read from
+the source directory, and checked for duplicates.  Once all file data has been
+written the completed inode, directory, fragment, export, uid/gid lookup and
+xattr tables are written.
+
+3.1 Compression options
+-----------------------
+
+Compressors can optionally support compression specific options (e.g.
+dictionary size).  If non-default compression options have been used, then
+these are stored here.
+
+3.2 Inodes
+----------
+
+Metadata (inodes and directories) are compressed in 8Kbyte blocks.  Each
+compressed block is prefixed by a two byte length, the top bit is set if the
+block is uncompressed.  A block will be uncompressed if the -noI option is set,
+or if the compressed block was larger than the uncompressed block.
+
+Inodes are packed into the metadata blocks, and are not aligned to block
+boundaries, therefore inodes overlap compressed blocks.  Inodes are identified
+by a 48-bit number which encodes the location of the compressed metadata block
+containing the inode, and the byte offset into that block where the inode is
+placed (<block, offset>).
+
+To maximise compression there are different inodes for each file type
+(regular file, directory, device, etc.), the inode contents and length
+varying with the type.
+
+To further maximise compression, two types of regular file inode and
+directory inode are defined: inodes optimised for frequently occurring
+regular files and directories, and extended types where extra
+information has to be stored.
+
+3.3 Directories
+---------------
+
+Like inodes, directories are packed into compressed metadata blocks, stored
+in a directory table.  Directories are accessed using the start address of
+the metablock containing the directory and the offset into the
+decompressed block (<block, offset>).
+
+Directories are organised in a slightly complex way, and are not simply
+a list of file names.  The organisation takes advantage of the
+fact that (in most cases) the inodes of the files will be in the same
+compressed metadata block, and therefore, can share the start block.
+Directories are therefore organised in a two level list, a directory
+header containing the shared start block value, and a sequence of directory
+entries, each of which share the shared start block.  A new directory header
+is written once/if the inode start block changes.  The directory
+header/directory entry list is repeated as many times as necessary.
+
+Directories are sorted, and can contain a directory index to speed up
+file lookup.  Directory indexes store one entry per metablock, each entry
+storing the index/filename mapping to the first directory header
+in each metadata block.  Directories are sorted in alphabetical order,
+and at lookup the index is scanned linearly looking for the first filename
+alphabetically larger than the filename being looked up.  At this point the
+location of the metadata block the filename is in has been found.
+The general idea of the index is to ensure only one metadata block needs to be
+decompressed to do a lookup irrespective of the length of the directory.
+This scheme has the advantage that it doesn't require extra memory overhead
+and doesn't require much extra storage on disk.
+
+3.4 File data
+-------------
+
+Regular files consist of a sequence of contiguous compressed blocks, and/or a
+compressed fragment block (tail-end packed block).   The compressed size
+of each datablock is stored in a block list contained within the
+file inode.
+
+To speed up access to datablocks when reading 'large' files (256 Mbytes or
+larger), the code implements an index cache that caches the mapping from
+block index to datablock location on disk.
+
+The index cache allows Squashfs to handle large files (up to 1.75 TiB) while
+retaining a simple and space-efficient block list on disk.  The cache
+is split into slots, caching up to eight 224 GiB files (128 KiB blocks).
+Larger files use multiple slots, with 1.75 TiB files using all 8 slots.
+The index cache is designed to be memory efficient, and by default uses
+16 KiB.
+
+3.5 Fragment lookup table
+-------------------------
+
+Regular files can contain a fragment index which is mapped to a fragment
+location on disk and compressed size using a fragment lookup table.  This
+fragment lookup table is itself stored compressed into metadata blocks.
+A second index table is used to locate these.  This second index table for
+speed of access (and because it is small) is read at mount time and cached
+in memory.
+
+3.6 Uid/gid lookup table
+------------------------
+
+For space efficiency regular files store uid and gid indexes, which are
+converted to 32-bit uids/gids using an id look up table.  This table is
+stored compressed into metadata blocks.  A second index table is used to
+locate these.  This second index table for speed of access (and because it
+is small) is read at mount time and cached in memory.
+
+3.7 Export table
+----------------
+
+To enable Squashfs filesystems to be exportable (via NFS etc.) filesystems
+can optionally (disabled with the -no-exports Mksquashfs option) contain
+an inode number to inode disk location lookup table.  This is required to
+enable Squashfs to map inode numbers passed in filehandles to the inode
+location on disk, which is necessary when the export code reinstantiates
+expired/flushed inodes.
+
+This table is stored compressed into metadata blocks.  A second index table is
+used to locate these.  This second index table for speed of access (and because
+it is small) is read at mount time and cached in memory.
+
+3.8 Xattr table
+---------------
+
+The xattr table contains extended attributes for each inode.  The xattrs
+for each inode are stored in a list, each list entry containing a type,
+name and value field.  The type field encodes the xattr prefix
+("user.", "trusted." etc) and it also encodes how the name/value fields
+should be interpreted.  Currently the type indicates whether the value
+is stored inline (in which case the value field contains the xattr value),
+or if it is stored out of line (in which case the value field stores a
+reference to where the actual value is stored).  This allows large values
+to be stored out of line improving scanning and lookup performance and it
+also allows values to be de-duplicated, the value being stored once, and
+all other occurrences holding an out of line reference to that value.
+
+The xattr lists are packed into compressed 8K metadata blocks.
+To reduce overhead in inodes, rather than storing the on-disk
+location of the xattr list inside each inode, a 32-bit xattr id
+is stored.  This xattr id is mapped into the location of the xattr
+list using a second xattr id lookup table.
+
+4. TODOS AND OUTSTANDING ISSUES
+-------------------------------
+
+4.1 Todo list
+-------------
+
+Implement ACL support.
+
+4.2 Squashfs internal cache
+---------------------------
+
+Blocks in Squashfs are compressed.  To avoid repeatedly decompressing
+recently accessed data Squashfs uses two small metadata and fragment caches.
+
+The cache is not used for file datablocks, these are decompressed and cached in
+the page-cache in the normal way.  The cache is used to temporarily cache
+fragment and metadata blocks which have been read as a result of a metadata
+(i.e. inode or directory) or fragment access.  Because metadata and fragments
+are packed together into blocks (to gain greater compression) the read of a
+particular piece of metadata or fragment will retrieve other metadata/fragments
+which have been packed with it, these because of locality-of-reference may be
+read in the near future. Temporarily caching them ensures they are available
+for near future access without requiring an additional read and decompress.
+
+In the future this internal cache may be replaced with an implementation which
+uses the kernel page cache.  Because the page cache operates on page sized
+units this may introduce additional complexity in terms of locking and
+associated race conditions.
diff --git a/Documentation/filesystems/sysfs-pci.txt b/Documentation/filesystems/sysfs-pci.txt
new file mode 100644
index 000000000..06f1d64c6
--- /dev/null
+++ b/Documentation/filesystems/sysfs-pci.txt
@@ -0,0 +1,131 @@
+Accessing PCI device resources through sysfs
+--------------------------------------------
+
+sysfs, usually mounted at /sys, provides access to PCI resources on platforms
+that support it.  For example, a given bus might look like this:
+
+     /sys/devices/pci0000:17
+     |-- 0000:17:00.0
+     |   |-- class
+     |   |-- config
+     |   |-- device
+     |   |-- enable
+     |   |-- irq
+     |   |-- local_cpus
+     |   |-- remove
+     |   |-- resource
+     |   |-- resource0
+     |   |-- resource1
+     |   |-- resource2
+     |   |-- revision
+     |   |-- rom
+     |   |-- subsystem_device
+     |   |-- subsystem_vendor
+     |   `-- vendor
+     `-- ...
+
+The topmost element describes the PCI domain and bus number.  In this case,
+the domain number is 0000 and the bus number is 17 (both values are in hex).
+This bus contains a single function device in slot 0.  The domain and bus
+numbers are reproduced for convenience.  Under the device directory are several
+files, each with their own function.
+
+       file		   function
+       ----		   --------
+       class		   PCI class (ascii, ro)
+       config		   PCI config space (binary, rw)
+       device		   PCI device (ascii, ro)
+       enable	           Whether the device is enabled (ascii, rw)
+       irq		   IRQ number (ascii, ro)
+       local_cpus	   nearby CPU mask (cpumask, ro)
+       remove		   remove device from kernel's list (ascii, wo)
+       resource		   PCI resource host addresses (ascii, ro)
+       resource0..N	   PCI resource N, if present (binary, mmap, rw[1])
+       resource0_wc..N_wc  PCI WC map resource N, if prefetchable (binary, mmap)
+       revision		   PCI revision (ascii, ro)
+       rom		   PCI ROM resource, if present (binary, ro)
+       subsystem_device	   PCI subsystem device (ascii, ro)
+       subsystem_vendor	   PCI subsystem vendor (ascii, ro)
+       vendor		   PCI vendor (ascii, ro)
+
+  ro - read only file
+  rw - file is readable and writable
+  wo - write only file
+  mmap - file is mmapable
+  ascii - file contains ascii text
+  binary - file contains binary data
+  cpumask - file contains a cpumask type
+
+[1] rw for RESOURCE_IO (I/O port) regions only
+
+The read only files are informational, writes to them will be ignored, with
+the exception of the 'rom' file.  Writable files can be used to perform
+actions on the device (e.g. changing config space, detaching a device).
+mmapable files are available via an mmap of the file at offset 0 and can be
+used to do actual device programming from userspace.  Note that some platforms
+don't support mmapping of certain resources, so be sure to check the return
+value from any attempted mmap.  The most notable of these are I/O port
+resources, which also provide read/write access.
+
+The 'enable' file provides a counter that indicates how many times the device 
+has been enabled.  If the 'enable' file currently returns '4', and a '1' is
+echoed into it, it will then return '5'.  Echoing a '0' into it will decrease
+the count.  Even when it returns to 0, though, some of the initialisation
+may not be reversed.  
+
+The 'rom' file is special in that it provides read-only access to the device's
+ROM file, if available.  It's disabled by default, however, so applications
+should write the string "1" to the file to enable it before attempting a read
+call, and disable it following the access by writing "0" to the file.  Note
+that the device must be enabled for a rom read to return data successfully.
+In the event a driver is not bound to the device, it can be enabled using the
+'enable' file, documented above.
+
+The 'remove' file is used to remove the PCI device, by writing a non-zero
+integer to the file.  This does not involve any kind of hot-plug functionality,
+e.g. powering off the device.  The device is removed from the kernel's list of
+PCI devices, the sysfs directory for it is removed, and the device will be
+removed from any drivers attached to it. Removal of PCI root buses is
+disallowed.
+
+Accessing legacy resources through sysfs
+----------------------------------------
+
+Legacy I/O port and ISA memory resources are also provided in sysfs if the
+underlying platform supports them.  They're located in the PCI class hierarchy,
+e.g.
+
+	/sys/class/pci_bus/0000:17/
+	|-- bridge -> ../../../devices/pci0000:17
+	|-- cpuaffinity
+	|-- legacy_io
+	`-- legacy_mem
+
+The legacy_io file is a read/write file that can be used by applications to
+do legacy port I/O.  The application should open the file, seek to the desired
+port (e.g. 0x3e8) and do a read or a write of 1, 2 or 4 bytes.  The legacy_mem
+file should be mmapped with an offset corresponding to the memory offset
+desired, e.g. 0xa0000 for the VGA frame buffer.  The application can then
+simply dereference the returned pointer (after checking for errors of course)
+to access legacy memory space.
+
+Supporting PCI access on new platforms
+--------------------------------------
+
+In order to support PCI resource mapping as described above, Linux platform
+code should ideally define ARCH_GENERIC_PCI_MMAP_RESOURCE and use the generic
+implementation of that functionality. To support the historical interface of
+mmap() through files in /proc/bus/pci, platforms may also set HAVE_PCI_MMAP.
+
+Alternatively, platforms which set HAVE_PCI_MMAP may provide their own
+implementation of pci_mmap_page_range() instead of defining
+ARCH_GENERIC_PCI_MMAP_RESOURCE.
+
+Platforms which support write-combining maps of PCI resources must define
+arch_can_pci_mmap_wc() which shall evaluate to non-zero at runtime when
+write-combining is permitted. Platforms which support maps of I/O resources
+define arch_can_pci_mmap_io() similarly.
+
+Legacy resources are protected by the HAVE_PCI_LEGACY define.  Platforms
+wishing to support legacy functionality should define it and provide
+pci_legacy_read, pci_legacy_write and pci_mmap_legacy_page_range functions.
diff --git a/Documentation/filesystems/sysfs-tagging.txt b/Documentation/filesystems/sysfs-tagging.txt
new file mode 100644
index 000000000..c7c8e6438
--- /dev/null
+++ b/Documentation/filesystems/sysfs-tagging.txt
@@ -0,0 +1,42 @@
+Sysfs tagging
+-------------
+
+(Taken almost verbatim from Eric Biederman's netns tagging patch
+commit msg)
+
+The problem.  Network devices show up in sysfs and with the network
+namespace active multiple devices with the same name can show up in
+the same directory, ouch!
+
+To avoid that problem and allow existing applications in network
+namespaces to see the same interface that is currently presented in
+sysfs, sysfs now has tagging directory support.
+
+By using the network namespace pointers as tags to separate out the
+the sysfs directory entries we ensure that we don't have conflicts
+in the directories and applications only see a limited set of
+the network devices.
+
+Each sysfs directory entry may be tagged with a namespace via the
+void *ns member of its kernfs_node.  If a directory entry is tagged,
+then kernfs_node->flags will have a flag between KOBJ_NS_TYPE_NONE
+and KOBJ_NS_TYPES, and ns will point to the namespace to which it
+belongs.
+
+Each sysfs superblock's kernfs_super_info contains an array void
+*ns[KOBJ_NS_TYPES].  When a task in a tagging namespace
+kobj_nstype first mounts sysfs, a new superblock is created.  It
+will be differentiated from other sysfs mounts by having its
+s_fs_info->ns[kobj_nstype] set to the new namespace.  Note that
+through bind mounting and mounts propagation, a task can easily view
+the contents of other namespaces' sysfs mounts.  Therefore, when a
+namespace exits, it will call kobj_ns_exit() to invalidate any
+kernfs_node->ns pointers pointing to it.
+
+Users of this interface:
+- define a type in the kobj_ns_type enumeration.
+- call kobj_ns_type_register() with its kobj_ns_type_operations which has
+  - current_ns() which returns current's namespace
+  - netlink_ns() which returns a socket's namespace
+  - initial_ns() which returns the initial namesapce
+- call kobj_ns_exit() when an individual tag is no longer valid
diff --git a/Documentation/filesystems/sysfs.txt b/Documentation/filesystems/sysfs.txt
new file mode 100644
index 000000000..2e38fafc1
--- /dev/null
+++ b/Documentation/filesystems/sysfs.txt
@@ -0,0 +1,383 @@
+
+sysfs - _The_ filesystem for exporting kernel objects. 
+
+Patrick Mochel	<mochel@osdl.org>
+Mike Murphy <mamurph@cs.clemson.edu>
+
+Revised:    16 August 2011
+Original:   10 January 2003
+
+
+What it is:
+~~~~~~~~~~~
+
+sysfs is a ram-based filesystem initially based on ramfs. It provides
+a means to export kernel data structures, their attributes, and the 
+linkages between them to userspace. 
+
+sysfs is tied inherently to the kobject infrastructure. Please read
+Documentation/kobject.txt for more information concerning the kobject
+interface. 
+
+
+Using sysfs
+~~~~~~~~~~~
+
+sysfs is always compiled in if CONFIG_SYSFS is defined. You can access
+it by doing:
+
+    mount -t sysfs sysfs /sys 
+
+
+Directory Creation
+~~~~~~~~~~~~~~~~~~
+
+For every kobject that is registered with the system, a directory is
+created for it in sysfs. That directory is created as a subdirectory
+of the kobject's parent, expressing internal object hierarchies to
+userspace. Top-level directories in sysfs represent the common
+ancestors of object hierarchies; i.e. the subsystems the objects
+belong to. 
+
+Sysfs internally stores a pointer to the kobject that implements a
+directory in the kernfs_node object associated with the directory. In
+the past this kobject pointer has been used by sysfs to do reference
+counting directly on the kobject whenever the file is opened or closed.
+With the current sysfs implementation the kobject reference count is
+only modified directly by the function sysfs_schedule_callback().
+
+
+Attributes
+~~~~~~~~~~
+
+Attributes can be exported for kobjects in the form of regular files in
+the filesystem. Sysfs forwards file I/O operations to methods defined
+for the attributes, providing a means to read and write kernel
+attributes.
+
+Attributes should be ASCII text files, preferably with only one value
+per file. It is noted that it may not be efficient to contain only one
+value per file, so it is socially acceptable to express an array of
+values of the same type. 
+
+Mixing types, expressing multiple lines of data, and doing fancy
+formatting of data is heavily frowned upon. Doing these things may get
+you publicly humiliated and your code rewritten without notice. 
+
+
+An attribute definition is simply:
+
+struct attribute {
+        char                    * name;
+        struct module		*owner;
+        umode_t                 mode;
+};
+
+
+int sysfs_create_file(struct kobject * kobj, const struct attribute * attr);
+void sysfs_remove_file(struct kobject * kobj, const struct attribute * attr);
+
+
+A bare attribute contains no means to read or write the value of the
+attribute. Subsystems are encouraged to define their own attribute
+structure and wrapper functions for adding and removing attributes for
+a specific object type. 
+
+For example, the driver model defines struct device_attribute like:
+
+struct device_attribute {
+	struct attribute	attr;
+	ssize_t (*show)(struct device *dev, struct device_attribute *attr,
+			char *buf);
+	ssize_t (*store)(struct device *dev, struct device_attribute *attr,
+			 const char *buf, size_t count);
+};
+
+int device_create_file(struct device *, const struct device_attribute *);
+void device_remove_file(struct device *, const struct device_attribute *);
+
+It also defines this helper for defining device attributes: 
+
+#define DEVICE_ATTR(_name, _mode, _show, _store) \
+struct device_attribute dev_attr_##_name = __ATTR(_name, _mode, _show, _store)
+
+For example, declaring
+
+static DEVICE_ATTR(foo, S_IWUSR | S_IRUGO, show_foo, store_foo);
+
+is equivalent to doing:
+
+static struct device_attribute dev_attr_foo = {
+	.attr = {
+		.name = "foo",
+		.mode = S_IWUSR | S_IRUGO,
+	},
+	.show = show_foo,
+	.store = store_foo,
+};
+
+
+Subsystem-Specific Callbacks
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+When a subsystem defines a new attribute type, it must implement a
+set of sysfs operations for forwarding read and write calls to the
+show and store methods of the attribute owners. 
+
+struct sysfs_ops {
+        ssize_t (*show)(struct kobject *, struct attribute *, char *);
+        ssize_t (*store)(struct kobject *, struct attribute *, const char *, size_t);
+};
+
+[ Subsystems should have already defined a struct kobj_type as a
+descriptor for this type, which is where the sysfs_ops pointer is
+stored. See the kobject documentation for more information. ]
+
+When a file is read or written, sysfs calls the appropriate method
+for the type. The method then translates the generic struct kobject
+and struct attribute pointers to the appropriate pointer types, and
+calls the associated methods. 
+
+
+To illustrate:
+
+#define to_dev(obj) container_of(obj, struct device, kobj)
+#define to_dev_attr(_attr) container_of(_attr, struct device_attribute, attr)
+
+static ssize_t dev_attr_show(struct kobject *kobj, struct attribute *attr,
+                             char *buf)
+{
+        struct device_attribute *dev_attr = to_dev_attr(attr);
+        struct device *dev = to_dev(kobj);
+        ssize_t ret = -EIO;
+
+        if (dev_attr->show)
+                ret = dev_attr->show(dev, dev_attr, buf);
+        if (ret >= (ssize_t)PAGE_SIZE) {
+                printk("dev_attr_show: %pS returned bad count\n",
+                                dev_attr->show);
+        }
+        return ret;
+}
+
+
+
+Reading/Writing Attribute Data
+~~~~~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+To read or write attributes, show() or store() methods must be
+specified when declaring the attribute. The method types should be as
+simple as those defined for device attributes:
+
+ssize_t (*show)(struct device *dev, struct device_attribute *attr, char *buf);
+ssize_t (*store)(struct device *dev, struct device_attribute *attr,
+                 const char *buf, size_t count);
+
+IOW, they should take only an object, an attribute, and a buffer as parameters.
+
+
+sysfs allocates a buffer of size (PAGE_SIZE) and passes it to the
+method. Sysfs will call the method exactly once for each read or
+write. This forces the following behavior on the method
+implementations: 
+
+- On read(2), the show() method should fill the entire buffer. 
+  Recall that an attribute should only be exporting one value, or an
+  array of similar values, so this shouldn't be that expensive. 
+
+  This allows userspace to do partial reads and forward seeks
+  arbitrarily over the entire file at will. If userspace seeks back to
+  zero or does a pread(2) with an offset of '0' the show() method will
+  be called again, rearmed, to fill the buffer.
+
+- On write(2), sysfs expects the entire buffer to be passed during the
+  first write. Sysfs then passes the entire buffer to the store() method.
+  A terminating null is added after the data on stores. This makes
+  functions like sysfs_streq() safe to use.
+
+  When writing sysfs files, userspace processes should first read the
+  entire file, modify the values it wishes to change, then write the
+  entire buffer back. 
+
+  Attribute method implementations should operate on an identical
+  buffer when reading and writing values. 
+
+Other notes:
+
+- Writing causes the show() method to be rearmed regardless of current
+  file position.
+
+- The buffer will always be PAGE_SIZE bytes in length. On i386, this
+  is 4096. 
+
+- show() methods should return the number of bytes printed into the
+  buffer.
+
+- show() should only use sysfs_emit() or sysfs_emit_at() when formatting
+  the value to be returned to user space.
+
+- store() should return the number of bytes used from the buffer. If the
+  entire buffer has been used, just return the count argument.
+
+- show() or store() can always return errors. If a bad value comes
+  through, be sure to return an error.
+
+- The object passed to the methods will be pinned in memory via sysfs
+  referencing counting its embedded object. However, the physical 
+  entity (e.g. device) the object represents may not be present. Be 
+  sure to have a way to check this, if necessary. 
+
+
+A very simple (and naive) implementation of a device attribute is:
+
+static ssize_t show_name(struct device *dev, struct device_attribute *attr,
+                         char *buf)
+{
+	return scnprintf(buf, PAGE_SIZE, "%s\n", dev->name);
+}
+
+static ssize_t store_name(struct device *dev, struct device_attribute *attr,
+                          const char *buf, size_t count)
+{
+        snprintf(dev->name, sizeof(dev->name), "%.*s",
+                 (int)min(count, sizeof(dev->name) - 1), buf);
+	return count;
+}
+
+static DEVICE_ATTR(name, S_IRUGO, show_name, store_name);
+
+
+(Note that the real implementation doesn't allow userspace to set the 
+name for a device.)
+
+
+Top Level Directory Layout
+~~~~~~~~~~~~~~~~~~~~~~~~~~
+
+The sysfs directory arrangement exposes the relationship of kernel
+data structures. 
+
+The top level sysfs directory looks like:
+
+block/
+bus/
+class/
+dev/
+devices/
+firmware/
+net/
+fs/
+
+devices/ contains a filesystem representation of the device tree. It maps
+directly to the internal kernel device tree, which is a hierarchy of
+struct device. 
+
+bus/ contains flat directory layout of the various bus types in the
+kernel. Each bus's directory contains two subdirectories:
+
+	devices/
+	drivers/
+
+devices/ contains symlinks for each device discovered in the system
+that point to the device's directory under root/.
+
+drivers/ contains a directory for each device driver that is loaded
+for devices on that particular bus (this assumes that drivers do not
+span multiple bus types).
+
+fs/ contains a directory for some filesystems.  Currently each
+filesystem wanting to export attributes must create its own hierarchy
+below fs/ (see ./fuse.txt for an example).
+
+dev/ contains two directories char/ and block/. Inside these two
+directories there are symlinks named <major>:<minor>.  These symlinks
+point to the sysfs directory for the given device.  /sys/dev provides a
+quick way to lookup the sysfs interface for a device from the result of
+a stat(2) operation.
+
+More information can driver-model specific features can be found in
+Documentation/driver-model/. 
+
+
+TODO: Finish this section.
+
+
+Current Interfaces
+~~~~~~~~~~~~~~~~~~
+
+The following interface layers currently exist in sysfs:
+
+
+- devices (include/linux/device.h)
+----------------------------------
+Structure:
+
+struct device_attribute {
+	struct attribute	attr;
+	ssize_t (*show)(struct device *dev, struct device_attribute *attr,
+			char *buf);
+	ssize_t (*store)(struct device *dev, struct device_attribute *attr,
+			 const char *buf, size_t count);
+};
+
+Declaring:
+
+DEVICE_ATTR(_name, _mode, _show, _store);
+
+Creation/Removal:
+
+int device_create_file(struct device *dev, const struct device_attribute * attr);
+void device_remove_file(struct device *dev, const struct device_attribute * attr);
+
+
+- bus drivers (include/linux/device.h)
+--------------------------------------
+Structure:
+
+struct bus_attribute {
+        struct attribute        attr;
+        ssize_t (*show)(struct bus_type *, char * buf);
+        ssize_t (*store)(struct bus_type *, const char * buf, size_t count);
+};
+
+Declaring:
+
+BUS_ATTR(_name, _mode, _show, _store)
+
+Creation/Removal:
+
+int bus_create_file(struct bus_type *, struct bus_attribute *);
+void bus_remove_file(struct bus_type *, struct bus_attribute *);
+
+
+- device drivers (include/linux/device.h)
+-----------------------------------------
+
+Structure:
+
+struct driver_attribute {
+        struct attribute        attr;
+        ssize_t (*show)(struct device_driver *, char * buf);
+        ssize_t (*store)(struct device_driver *, const char * buf,
+                         size_t count);
+};
+
+Declaring:
+
+DRIVER_ATTR_RO(_name)
+DRIVER_ATTR_RW(_name)
+
+Creation/Removal:
+
+int driver_create_file(struct device_driver *, const struct driver_attribute *);
+void driver_remove_file(struct device_driver *, const struct driver_attribute *);
+
+
+Documentation
+~~~~~~~~~~~~~
+
+The sysfs directory structure and the attributes in each directory define an
+ABI between the kernel and user space. As for any ABI, it is important that
+this ABI is stable and properly documented. All new sysfs attributes must be
+documented in Documentation/ABI. See also Documentation/ABI/README for more
+information.
diff --git a/Documentation/filesystems/sysv-fs.txt b/Documentation/filesystems/sysv-fs.txt
new file mode 100644
index 000000000..253b50d13
--- /dev/null
+++ b/Documentation/filesystems/sysv-fs.txt
@@ -0,0 +1,197 @@
+It implements all of
+  - Xenix FS,
+  - SystemV/386 FS,
+  - Coherent FS.
+
+To install:
+* Answer the 'System V and Coherent filesystem support' question with 'y'
+  when configuring the kernel.
+* To mount a disk or a partition, use
+    mount [-r] -t sysv device mountpoint
+  The file system type names
+               -t sysv
+               -t xenix
+               -t coherent
+  may be used interchangeably, but the last two will eventually disappear.
+
+Bugs in the present implementation:
+- Coherent FS:
+  - The "free list interleave" n:m is currently ignored.
+  - Only file systems with no filesystem name and no pack name are recognized.
+  (See Coherent "man mkfs" for a description of these features.)
+- SystemV Release 2 FS:
+  The superblock is only searched in the blocks 9, 15, 18, which
+  corresponds to the beginning of track 1 on floppy disks. No support
+  for this FS on hard disk yet.
+
+
+These filesystems are rather similar. Here is a comparison with Minix FS:
+
+* Linux fdisk reports on partitions
+  - Minix FS     0x81 Linux/Minix
+  - Xenix FS     ??
+  - SystemV FS   ??
+  - Coherent FS  0x08 AIX bootable
+
+* Size of a block or zone (data allocation unit on disk)
+  - Minix FS     1024
+  - Xenix FS     1024 (also 512 ??)
+  - SystemV FS   1024 (also 512 and 2048)
+  - Coherent FS   512
+
+* General layout: all have one boot block, one super block and
+  separate areas for inodes and for directories/data.
+  On SystemV Release 2 FS (e.g. Microport) the first track is reserved and
+  all the block numbers (including the super block) are offset by one track.
+
+* Byte ordering of "short" (16 bit entities) on disk:
+  - Minix FS     little endian  0 1
+  - Xenix FS     little endian  0 1
+  - SystemV FS   little endian  0 1
+  - Coherent FS  little endian  0 1
+  Of course, this affects only the file system, not the data of files on it!
+
+* Byte ordering of "long" (32 bit entities) on disk:
+  - Minix FS     little endian  0 1 2 3
+  - Xenix FS     little endian  0 1 2 3
+  - SystemV FS   little endian  0 1 2 3
+  - Coherent FS  PDP-11         2 3 0 1
+  Of course, this affects only the file system, not the data of files on it!
+
+* Inode on disk: "short", 0 means non-existent, the root dir ino is:
+  - Minix FS                            1
+  - Xenix FS, SystemV FS, Coherent FS   2
+
+* Maximum number of hard links to a file:
+  - Minix FS     250
+  - Xenix FS     ??
+  - SystemV FS   ??
+  - Coherent FS  >=10000
+
+* Free inode management:
+  - Minix FS                             a bitmap
+  - Xenix FS, SystemV FS, Coherent FS
+      There is a cache of a certain number of free inodes in the super-block.
+      When it is exhausted, new free inodes are found using a linear search.
+
+* Free block management:
+  - Minix FS                             a bitmap
+  - Xenix FS, SystemV FS, Coherent FS
+      Free blocks are organized in a "free list". Maybe a misleading term,
+      since it is not true that every free block contains a pointer to
+      the next free block. Rather, the free blocks are organized in chunks
+      of limited size, and every now and then a free block contains pointers
+      to the free blocks pertaining to the next chunk; the first of these
+      contains pointers and so on. The list terminates with a "block number"
+      0 on Xenix FS and SystemV FS, with a block zeroed out on Coherent FS.
+
+* Super-block location:
+  - Minix FS     block 1 = bytes 1024..2047
+  - Xenix FS     block 1 = bytes 1024..2047
+  - SystemV FS   bytes 512..1023
+  - Coherent FS  block 1 = bytes 512..1023
+
+* Super-block layout:
+  - Minix FS
+                    unsigned short s_ninodes;
+                    unsigned short s_nzones;
+                    unsigned short s_imap_blocks;
+                    unsigned short s_zmap_blocks;
+                    unsigned short s_firstdatazone;
+                    unsigned short s_log_zone_size;
+                    unsigned long s_max_size;
+                    unsigned short s_magic;
+  - Xenix FS, SystemV FS, Coherent FS
+                    unsigned short s_firstdatazone;
+                    unsigned long  s_nzones;
+                    unsigned short s_fzone_count;
+                    unsigned long  s_fzones[NICFREE];
+                    unsigned short s_finode_count;
+                    unsigned short s_finodes[NICINOD];
+                    char           s_flock;
+                    char           s_ilock;
+                    char           s_modified;
+                    char           s_rdonly;
+                    unsigned long  s_time;
+                    short          s_dinfo[4]; -- SystemV FS only
+                    unsigned long  s_free_zones;
+                    unsigned short s_free_inodes;
+                    short          s_dinfo[4]; -- Xenix FS only
+                    unsigned short s_interleave_m,s_interleave_n; -- Coherent FS only
+                    char           s_fname[6];
+                    char           s_fpack[6];
+    then they differ considerably:
+        Xenix FS
+                    char           s_clean;
+                    char           s_fill[371];
+                    long           s_magic;
+                    long           s_type;
+        SystemV FS
+                    long           s_fill[12 or 14];
+                    long           s_state;
+                    long           s_magic;
+                    long           s_type;
+        Coherent FS
+                    unsigned long  s_unique;
+    Note that Coherent FS has no magic.
+
+* Inode layout:
+  - Minix FS
+                    unsigned short i_mode;
+                    unsigned short i_uid;
+                    unsigned long  i_size;
+                    unsigned long  i_time;
+                    unsigned char  i_gid;
+                    unsigned char  i_nlinks;
+                    unsigned short i_zone[7+1+1];
+  - Xenix FS, SystemV FS, Coherent FS
+                    unsigned short i_mode;
+                    unsigned short i_nlink;
+                    unsigned short i_uid;
+                    unsigned short i_gid;
+                    unsigned long  i_size;
+                    unsigned char  i_zone[3*(10+1+1+1)];
+                    unsigned long  i_atime;
+                    unsigned long  i_mtime;
+                    unsigned long  i_ctime;
+
+* Regular file data blocks are organized as
+  - Minix FS
+               7 direct blocks
+               1 indirect block (pointers to blocks)
+               1 double-indirect block (pointer to pointers to blocks)
+  - Xenix FS, SystemV FS, Coherent FS
+              10 direct blocks
+               1 indirect block (pointers to blocks)
+               1 double-indirect block (pointer to pointers to blocks)
+               1 triple-indirect block (pointer to pointers to pointers to blocks)
+
+* Inode size, inodes per block
+  - Minix FS        32   32
+  - Xenix FS        64   16
+  - SystemV FS      64   16
+  - Coherent FS     64    8
+
+* Directory entry on disk
+  - Minix FS
+                    unsigned short inode;
+                    char name[14/30];
+  - Xenix FS, SystemV FS, Coherent FS
+                    unsigned short inode;
+                    char name[14];
+
+* Dir entry size, dir entries per block
+  - Minix FS     16/32    64/32
+  - Xenix FS     16       64
+  - SystemV FS   16       64
+  - Coherent FS  16       32
+
+* How to implement symbolic links such that the host fsck doesn't scream:
+  - Minix FS     normal
+  - Xenix FS     kludge: as regular files with  chmod 1000
+  - SystemV FS   ??
+  - Coherent FS  kludge: as regular files with  chmod 1000
+
+
+Notation: We often speak of a "block" but mean a zone (the allocation unit)
+and not the disk driver's notion of "block".
diff --git a/Documentation/filesystems/tmpfs.txt b/Documentation/filesystems/tmpfs.txt
new file mode 100644
index 000000000..d06e9a59a
--- /dev/null
+++ b/Documentation/filesystems/tmpfs.txt
@@ -0,0 +1,149 @@
+Tmpfs is a file system which keeps all files in virtual memory.
+
+
+Everything in tmpfs is temporary in the sense that no files will be
+created on your hard drive. If you unmount a tmpfs instance,
+everything stored therein is lost.
+
+tmpfs puts everything into the kernel internal caches and grows and
+shrinks to accommodate the files it contains and is able to swap
+unneeded pages out to swap space. It has maximum size limits which can
+be adjusted on the fly via 'mount -o remount ...'
+
+If you compare it to ramfs (which was the template to create tmpfs)
+you gain swapping and limit checking. Another similar thing is the RAM
+disk (/dev/ram*), which simulates a fixed size hard disk in physical
+RAM, where you have to create an ordinary filesystem on top. Ramdisks
+cannot swap and you do not have the possibility to resize them. 
+
+Since tmpfs lives completely in the page cache and on swap, all tmpfs
+pages will be shown as "Shmem" in /proc/meminfo and "Shared" in
+free(1). Notice that these counters also include shared memory
+(shmem, see ipcs(1)). The most reliable way to get the count is
+using df(1) and du(1).
+
+tmpfs has the following uses:
+
+1) There is always a kernel internal mount which you will not see at
+   all. This is used for shared anonymous mappings and SYSV shared
+   memory. 
+
+   This mount does not depend on CONFIG_TMPFS. If CONFIG_TMPFS is not
+   set, the user visible part of tmpfs is not build. But the internal
+   mechanisms are always present.
+
+2) glibc 2.2 and above expects tmpfs to be mounted at /dev/shm for
+   POSIX shared memory (shm_open, shm_unlink). Adding the following
+   line to /etc/fstab should take care of this:
+
+	tmpfs	/dev/shm	tmpfs	defaults	0 0
+
+   Remember to create the directory that you intend to mount tmpfs on
+   if necessary.
+
+   This mount is _not_ needed for SYSV shared memory. The internal
+   mount is used for that. (In the 2.3 kernel versions it was
+   necessary to mount the predecessor of tmpfs (shm fs) to use SYSV
+   shared memory)
+
+3) Some people (including me) find it very convenient to mount it
+   e.g. on /tmp and /var/tmp and have a big swap partition. And now
+   loop mounts of tmpfs files do work, so mkinitrd shipped by most
+   distributions should succeed with a tmpfs /tmp.
+
+4) And probably a lot more I do not know about :-)
+
+
+tmpfs has three mount options for sizing:
+
+size:      The limit of allocated bytes for this tmpfs instance. The 
+           default is half of your physical RAM without swap. If you
+           oversize your tmpfs instances the machine will deadlock
+           since the OOM handler will not be able to free that memory.
+nr_blocks: The same as size, but in blocks of PAGE_SIZE.
+nr_inodes: The maximum number of inodes for this instance. The default
+           is half of the number of your physical RAM pages, or (on a
+           machine with highmem) the number of lowmem RAM pages,
+           whichever is the lower.
+
+These parameters accept a suffix k, m or g for kilo, mega and giga and
+can be changed on remount.  The size parameter also accepts a suffix %
+to limit this tmpfs instance to that percentage of your physical RAM:
+the default, when neither size nor nr_blocks is specified, is size=50%
+
+If nr_blocks=0 (or size=0), blocks will not be limited in that instance;
+if nr_inodes=0, inodes will not be limited.  It is generally unwise to
+mount with such options, since it allows any user with write access to
+use up all the memory on the machine; but enhances the scalability of
+that instance in a system with many cpus making intensive use of it.
+
+
+tmpfs has a mount option to set the NUMA memory allocation policy for
+all files in that instance (if CONFIG_NUMA is enabled) - which can be
+adjusted on the fly via 'mount -o remount ...'
+
+mpol=default             use the process allocation policy
+                         (see set_mempolicy(2))
+mpol=prefer:Node         prefers to allocate memory from the given Node
+mpol=bind:NodeList       allocates memory only from nodes in NodeList
+mpol=interleave          prefers to allocate from each node in turn
+mpol=interleave:NodeList allocates from each node of NodeList in turn
+mpol=local		 prefers to allocate memory from the local node
+
+NodeList format is a comma-separated list of decimal numbers and ranges,
+a range being two hyphen-separated decimal numbers, the smallest and
+largest node numbers in the range.  For example, mpol=bind:0-3,5,7,9-15
+
+A memory policy with a valid NodeList will be saved, as specified, for
+use at file creation time.  When a task allocates a file in the file
+system, the mount option memory policy will be applied with a NodeList,
+if any, modified by the calling task's cpuset constraints
+[See Documentation/cgroup-v1/cpusets.txt] and any optional flags, listed
+below.  If the resulting NodeLists is the empty set, the effective memory
+policy for the file will revert to "default" policy.
+
+NUMA memory allocation policies have optional flags that can be used in
+conjunction with their modes.  These optional flags can be specified
+when tmpfs is mounted by appending them to the mode before the NodeList.
+See Documentation/admin-guide/mm/numa_memory_policy.rst for a list of
+all available memory allocation policy mode flags and their effect on
+memory policy.
+
+	=static		is equivalent to	MPOL_F_STATIC_NODES
+	=relative	is equivalent to	MPOL_F_RELATIVE_NODES
+
+For example, mpol=bind=static:NodeList, is the equivalent of an
+allocation policy of MPOL_BIND | MPOL_F_STATIC_NODES.
+
+Note that trying to mount a tmpfs with an mpol option will fail if the
+running kernel does not support NUMA; and will fail if its nodelist
+specifies a node which is not online.  If your system relies on that
+tmpfs being mounted, but from time to time runs a kernel built without
+NUMA capability (perhaps a safe recovery kernel), or with fewer nodes
+online, then it is advisable to omit the mpol option from automatic
+mount options.  It can be added later, when the tmpfs is already mounted
+on MountPoint, by 'mount -o remount,mpol=Policy:NodeList MountPoint'.
+
+
+To specify the initial root directory you can use the following mount
+options:
+
+mode:	The permissions as an octal number
+uid:	The user id 
+gid:	The group id
+
+These options do not have any effect on remount. You can change these
+parameters with chmod(1), chown(1) and chgrp(1) on a mounted filesystem.
+
+
+So 'mount -t tmpfs -o size=10G,nr_inodes=10k,mode=700 tmpfs /mytmpfs'
+will give you tmpfs instance on /mytmpfs which can allocate 10GB
+RAM/SWAP in 10240 inodes and it is only accessible by root.
+
+
+Author:
+   Christoph Rohland <cr@sap.com>, 1.12.01
+Updated:
+   Hugh Dickins, 4 June 2007
+Updated:
+   KOSAKI Motohiro, 16 Mar 2010
diff --git a/Documentation/filesystems/ubifs.txt b/Documentation/filesystems/ubifs.txt
new file mode 100644
index 000000000..a0a61d2f3
--- /dev/null
+++ b/Documentation/filesystems/ubifs.txt
@@ -0,0 +1,119 @@
+Introduction
+=============
+
+UBIFS file-system stands for UBI File System. UBI stands for "Unsorted
+Block Images". UBIFS is a flash file system, which means it is designed
+to work with flash devices. It is important to understand, that UBIFS
+is completely different to any traditional file-system in Linux, like
+Ext2, XFS, JFS, etc. UBIFS represents a separate class of file-systems
+which work with MTD devices, not block devices. The other Linux
+file-system of this class is JFFS2.
+
+To make it more clear, here is a small comparison of MTD devices and
+block devices.
+
+1 MTD devices represent flash devices and they consist of eraseblocks of
+  rather large size, typically about 128KiB. Block devices consist of
+  small blocks, typically 512 bytes.
+2 MTD devices support 3 main operations - read from some offset within an
+  eraseblock, write to some offset within an eraseblock, and erase a whole
+  eraseblock. Block  devices support 2 main operations - read a whole
+  block and write a whole block.
+3 The whole eraseblock has to be erased before it becomes possible to
+  re-write its contents. Blocks may be just re-written.
+4 Eraseblocks become worn out after some number of erase cycles -
+  typically 100K-1G for SLC NAND and NOR flashes, and 1K-10K for MLC
+  NAND flashes. Blocks do not have the wear-out property.
+5 Eraseblocks may become bad (only on NAND flashes) and software should
+  deal with this. Blocks on hard drives typically do not become bad,
+  because hardware has mechanisms to substitute bad blocks, at least in
+  modern LBA disks.
+
+It should be quite obvious why UBIFS is very different to traditional
+file-systems.
+
+UBIFS works on top of UBI. UBI is a separate software layer which may be
+found in drivers/mtd/ubi. UBI is basically a volume management and
+wear-leveling layer. It provides so called UBI volumes which is a higher
+level abstraction than a MTD device. The programming model of UBI devices
+is very similar to MTD devices - they still consist of large eraseblocks,
+they have read/write/erase operations, but UBI devices are devoid of
+limitations like wear and bad blocks (items 4 and 5 in the above list).
+
+In a sense, UBIFS is a next generation of JFFS2 file-system, but it is
+very different and incompatible to JFFS2. The following are the main
+differences.
+
+* JFFS2 works on top of MTD devices, UBIFS depends on UBI and works on
+  top of UBI volumes.
+* JFFS2 does not have on-media index and has to build it while mounting,
+  which requires full media scan. UBIFS maintains the FS indexing
+  information on the flash media and does not require full media scan,
+  so it mounts many times faster than JFFS2.
+* JFFS2 is a write-through file-system, while UBIFS supports write-back,
+  which makes UBIFS much faster on writes.
+
+Similarly to JFFS2, UBIFS supports on-the-flight compression which makes
+it possible to fit quite a lot of data to the flash.
+
+Similarly to JFFS2, UBIFS is tolerant of unclean reboots and power-cuts.
+It does not need stuff like fsck.ext2. UBIFS automatically replays its
+journal and recovers from crashes, ensuring that the on-flash data
+structures are consistent.
+
+UBIFS scales logarithmically (most of the data structures it uses are
+trees), so the mount time and memory consumption do not linearly depend
+on the flash size, like in case of JFFS2. This is because UBIFS
+maintains the FS index on the flash media. However, UBIFS depends on
+UBI, which scales linearly. So overall UBI/UBIFS stack scales linearly.
+Nevertheless, UBI/UBIFS scales considerably better than JFFS2.
+
+The authors of UBIFS believe, that it is possible to develop UBI2 which
+would scale logarithmically as well. UBI2 would support the same API as UBI,
+but it would be binary incompatible to UBI. So UBIFS would not need to be
+changed to use UBI2
+
+
+Mount options
+=============
+
+(*) == default.
+
+bulk_read		read more in one go to take advantage of flash
+			media that read faster sequentially
+no_bulk_read (*)	do not bulk-read
+no_chk_data_crc (*)	skip checking of CRCs on data nodes in order to
+			improve read performance. Use this option only
+			if the flash media is highly reliable. The effect
+			of this option is that corruption of the contents
+			of a file can go unnoticed.
+chk_data_crc		do not skip checking CRCs on data nodes
+compr=none              override default compressor and set it to "none"
+compr=lzo               override default compressor and set it to "lzo"
+compr=zlib              override default compressor and set it to "zlib"
+
+
+Quick usage instructions
+========================
+
+The UBI volume to mount is specified using "ubiX_Y" or "ubiX:NAME" syntax,
+where "X" is UBI device number, "Y" is UBI volume number, and "NAME" is
+UBI volume name.
+
+Mount volume 0 on UBI device 0 to /mnt/ubifs:
+$ mount -t ubifs ubi0_0 /mnt/ubifs
+
+Mount "rootfs" volume of UBI device 0 to /mnt/ubifs ("rootfs" is volume
+name):
+$ mount -t ubifs ubi0:rootfs /mnt/ubifs
+
+The following is an example of the kernel boot arguments to attach mtd0
+to UBI and mount volume "rootfs":
+ubi.mtd=0 root=ubi0:rootfs rootfstype=ubifs
+
+References
+==========
+
+UBIFS documentation and FAQ/HOWTO at the MTD web site:
+http://www.linux-mtd.infradead.org/doc/ubifs.html
+http://www.linux-mtd.infradead.org/faq/ubifs.html
diff --git a/Documentation/filesystems/udf.txt b/Documentation/filesystems/udf.txt
new file mode 100644
index 000000000..e2f2faf32
--- /dev/null
+++ b/Documentation/filesystems/udf.txt
@@ -0,0 +1,66 @@
+*
+* Documentation/filesystems/udf.txt
+*
+
+If you encounter problems with reading UDF discs using this driver,
+please report them according to MAINTAINERS file.
+
+Write support requires a block driver which supports writing.  Currently
+dvd+rw drives and media support true random sector writes, and so a udf
+filesystem on such devices can be directly mounted read/write.  CD-RW
+media however, does not support this.  Instead the media can be formatted
+for packet mode using the utility cdrwtool, then the pktcdvd driver can
+be bound to the underlying cd device to provide the required buffering
+and read-modify-write cycles to allow the filesystem random sector writes
+while providing the hardware with only full packet writes.  While not
+required for dvd+rw media, use of the pktcdvd driver often enhances
+performance due to very poor read-modify-write support supplied internally
+by drive firmware.
+
+-------------------------------------------------------------------------------
+The following mount options are supported:
+
+	gid=		Set the default group.
+	umask=		Set the default umask.
+	mode=		Set the default file permissions.
+	dmode=		Set the default directory permissions.
+	uid=		Set the default user.
+	bs=		Set the block size.
+	unhide		Show otherwise hidden files.
+	undelete	Show deleted files in lists.
+	adinicb		Embed data in the inode (default)
+	noadinicb	Don't embed data in the inode
+	shortad		Use short ad's
+	longad		Use long ad's (default)
+	nostrict	Unset strict conformance
+	iocharset=	Set the NLS character set
+
+The uid= and gid= options need a bit more explaining.  They will accept a
+decimal numeric value and all inodes on that mount will then appear as
+belonging to that uid and gid.  Mount options also accept the string "forget".
+The forget option causes all IDs to be written to disk as -1 which is a way
+of UDF standard to indicate that IDs are not supported for these files .
+
+For typical desktop use of removable media, you should set the ID to that of
+the interactively logged on user, and also specify the forget option.  This way
+the interactive user will always see the files on the disk as belonging to him.
+
+The remaining are for debugging and disaster recovery:
+
+	novrs		Skip volume sequence recognition 
+
+The following expect a offset from 0.
+
+	session=	Set the CDROM session (default= last session)
+	anchor=		Override standard anchor location. (default= 256)
+	lastblock=	Set the last block of the filesystem/
+
+-------------------------------------------------------------------------------
+
+
+For the latest version and toolset see:
+	https://github.com/pali/udftools
+
+Documentation on UDF and ECMA 167 is available FREE from:
+	http://www.osta.org/
+	http://www.ecma-international.org/
diff --git a/Documentation/filesystems/ufs.txt b/Documentation/filesystems/ufs.txt
new file mode 100644
index 000000000..7a602adec
--- /dev/null
+++ b/Documentation/filesystems/ufs.txt
@@ -0,0 +1,60 @@
+USING UFS
+=========
+
+mount -t ufs -o ufstype=type_of_ufs device dir
+
+
+UFS OPTIONS
+===========
+
+ufstype=type_of_ufs
+	UFS is a file system widely used in different operating systems.
+	The problem are differences among implementations. Features of
+	some implementations are undocumented, so its hard to recognize
+	type of ufs automatically. That's why user must specify type of 
+	ufs manually by mount option ufstype. Possible values are:
+
+	old	old format of ufs
+		default value, supported as read-only
+
+	44bsd	used in FreeBSD, NetBSD, OpenBSD
+		supported as read-write
+
+	ufs2    used in FreeBSD 5.x
+		supported as read-write
+
+	5xbsd	synonym for ufs2
+
+	sun	used in SunOS (Solaris)
+		supported as read-write
+
+	sunx86	used in SunOS for Intel (Solarisx86)
+		supported as read-write
+
+	hp	used in HP-UX
+		supported as read-only
+
+	nextstep
+		used in NextStep
+		supported as read-only
+
+	nextstep-cd
+		used for NextStep CDROMs (block_size == 2048)
+		supported as read-only
+
+	openstep
+		used in OpenStep
+		supported as read-only
+
+
+POSSIBLE PROBLEMS
+=================
+
+See next section, if you have any.
+
+
+BUG REPORTS
+===========
+
+Any ufs bug report you can send to daniel.pirkl@email.cz or
+to dushistov@mail.ru (do not send partition tables bug reports).
diff --git a/Documentation/filesystems/vfat.txt b/Documentation/filesystems/vfat.txt
new file mode 100644
index 000000000..91031298b
--- /dev/null
+++ b/Documentation/filesystems/vfat.txt
@@ -0,0 +1,347 @@
+USING VFAT
+----------------------------------------------------------------------
+To use the vfat filesystem, use the filesystem type 'vfat'.  i.e.
+  mount -t vfat /dev/fd0 /mnt
+
+No special partition formatter is required.  mkdosfs will work fine
+if you want to format from within Linux.
+
+VFAT MOUNT OPTIONS
+----------------------------------------------------------------------
+uid=###       -- Set the owner of all files on this filesystem.
+		 The default is the uid of current process.
+
+gid=###       -- Set the group of all files on this filesystem.
+		 The default is the gid of current process.
+
+umask=###     -- The permission mask (for files and directories, see umask(1)).
+                 The default is the umask of current process.
+
+dmask=###     -- The permission mask for the directory.
+                 The default is the umask of current process.
+
+fmask=###     -- The permission mask for files.
+                 The default is the umask of current process.
+
+allow_utime=### -- This option controls the permission check of mtime/atime.
+
+                  20 - If current process is in group of file's group ID,
+                       you can change timestamp.
+                   2 - Other users can change timestamp.
+
+                 The default is set from `dmask' option. (If the directory is
+                 writable, utime(2) is also allowed. I.e. ~dmask & 022)
+
+                 Normally utime(2) checks current process is owner of
+                 the file, or it has CAP_FOWNER capability.  But FAT
+                 filesystem doesn't have uid/gid on disk, so normal
+                 check is too unflexible. With this option you can
+                 relax it.
+
+codepage=###  -- Sets the codepage number for converting to shortname
+		 characters on FAT filesystem.
+		 By default, FAT_DEFAULT_CODEPAGE setting is used.
+
+iocharset=<name> -- Character set to use for converting between the
+		 encoding is used for user visible filename and 16 bit
+		 Unicode characters. Long filenames are stored on disk
+		 in Unicode format, but Unix for the most part doesn't
+		 know how to deal with Unicode.
+		 By default, FAT_DEFAULT_IOCHARSET setting is used.
+
+		 There is also an option of doing UTF-8 translations
+		 with the utf8 option.
+
+		 NOTE: "iocharset=utf8" is not recommended. If unsure,
+		 you should consider the following option instead.
+
+utf8=<bool>   -- UTF-8 is the filesystem safe version of Unicode that
+		 is used by the console. It can be enabled or disabled
+		 for the filesystem with this option.
+		 If 'uni_xlate' gets set, UTF-8 gets disabled.
+		 By default, FAT_DEFAULT_UTF8 setting is used.
+
+uni_xlate=<bool> -- Translate unhandled Unicode characters to special
+		 escaped sequences.  This would let you backup and
+		 restore filenames that are created with any Unicode
+		 characters.  Until Linux supports Unicode for real,
+		 this gives you an alternative.  Without this option,
+		 a '?' is used when no translation is possible.  The
+		 escape character is ':' because it is otherwise
+		 illegal on the vfat filesystem.  The escape sequence
+		 that gets used is ':' and the four digits of hexadecimal
+		 unicode.
+
+nonumtail=<bool> -- When creating 8.3 aliases, normally the alias will
+                 end in '~1' or tilde followed by some number.  If this
+                 option is set, then if the filename is 
+                 "longfilename.txt" and "longfile.txt" does not
+                 currently exist in the directory, 'longfile.txt' will
+                 be the short alias instead of 'longfi~1.txt'. 
+                  
+usefree       -- Use the "free clusters" value stored on FSINFO. It'll
+                 be used to determine number of free clusters without
+                 scanning disk. But it's not used by default, because
+                 recent Windows don't update it correctly in some
+                 case. If you are sure the "free clusters" on FSINFO is
+                 correct, by this option you can avoid scanning disk.
+
+quiet         -- Stops printing certain warning messages.
+
+check=s|r|n   -- Case sensitivity checking setting.
+                 s: strict, case sensitive
+                 r: relaxed, case insensitive
+                 n: normal, default setting, currently case insensitive
+
+nocase        -- This was deprecated for vfat. Use shortname=win95 instead.
+
+shortname=lower|win95|winnt|mixed
+	      -- Shortname display/create setting.
+		 lower: convert to lowercase for display,
+			emulate the Windows 95 rule for create.
+		 win95: emulate the Windows 95 rule for display/create.
+		 winnt: emulate the Windows NT rule for display/create.
+		 mixed: emulate the Windows NT rule for display,
+			emulate the Windows 95 rule for create.
+		 Default setting is `mixed'.
+
+tz=UTC        -- Interpret timestamps as UTC rather than local time.
+                 This option disables the conversion of timestamps
+                 between local time (as used by Windows on FAT) and UTC
+                 (which Linux uses internally).  This is particularly
+                 useful when mounting devices (like digital cameras)
+                 that are set to UTC in order to avoid the pitfalls of
+                 local time.
+time_offset=minutes
+	      -- Set offset for conversion of timestamps from local time
+		 used by FAT to UTC. I.e. <minutes> minutes will be subtracted
+		 from each timestamp to convert it to UTC used internally by
+		 Linux. This is useful when time zone set in sys_tz is
+		 not the time zone used by the filesystem. Note that this
+		 option still does not provide correct time stamps in all
+		 cases in presence of DST - time stamps in a different DST
+		 setting will be off by one hour.
+
+showexec      -- If set, the execute permission bits of the file will be
+		 allowed only if the extension part of the name is .EXE,
+		 .COM, or .BAT. Not set by default.
+
+debug         -- Can be set, but unused by the current implementation.
+
+sys_immutable -- If set, ATTR_SYS attribute on FAT is handled as
+		 IMMUTABLE flag on Linux. Not set by default.
+
+flush         -- If set, the filesystem will try to flush to disk more
+		 early than normal. Not set by default.
+
+rodir	      -- FAT has the ATTR_RO (read-only) attribute. On Windows,
+		 the ATTR_RO of the directory will just be ignored,
+		 and is used only by applications as a flag (e.g. it's set
+		 for the customized folder).
+
+		 If you want to use ATTR_RO as read-only flag even for
+		 the directory, set this option.
+
+errors=panic|continue|remount-ro
+	      -- specify FAT behavior on critical errors: panic, continue
+		 without doing anything or remount the partition in
+		 read-only mode (default behavior).
+
+discard       -- If set, issues discard/TRIM commands to the block
+		 device when blocks are freed. This is useful for SSD devices
+		 and sparse/thinly-provisoned LUNs.
+
+nfs=stale_rw|nostale_ro
+		Enable this only if you want to export the FAT filesystem
+		over NFS.
+
+		stale_rw: This option maintains an index (cache) of directory
+		inodes by i_logstart which is used by the nfs-related code to
+		improve look-ups. Full file operations (read/write) over NFS is
+		supported but with cache eviction at NFS server, this could
+		result in ESTALE issues.
+
+		nostale_ro: This option bases the inode number and filehandle
+		on the on-disk location of a file in the MS-DOS directory entry.
+		This ensures that ESTALE will not be returned after a file is
+		evicted from the inode cache. However, it means that operations
+		such as rename, create and unlink could cause filehandles that
+		previously pointed at one file to point at a different file,
+		potentially causing data corruption. For this reason, this
+		option also mounts the filesystem readonly.
+
+		To maintain backward compatibility, '-o nfs' is also accepted,
+		defaulting to stale_rw
+
+dos1xfloppy  -- If set, use a fallback default BIOS Parameter Block
+		configuration, determined by backing device size. These static
+		parameters match defaults assumed by DOS 1.x for 160 kiB,
+		180 kiB, 320 kiB, and 360 kiB floppies and floppy images.
+
+
+<bool>: 0,1,yes,no,true,false
+
+LIMITATION
+---------------------------------------------------------------------
+* The fallocated region of file is discarded at umount/evict time
+  when using fallocate with FALLOC_FL_KEEP_SIZE.
+  So, User should assume that fallocated region can be discarded at
+  last close if there is memory pressure resulting in eviction of
+  the inode from the memory. As a result, for any dependency on
+  the fallocated region, user should make sure to recheck fallocate
+  after reopening the file.
+
+TODO
+----------------------------------------------------------------------
+* Need to get rid of the raw scanning stuff.  Instead, always use
+  a get next directory entry approach.  The only thing left that uses
+  raw scanning is the directory renaming code.
+
+
+POSSIBLE PROBLEMS
+----------------------------------------------------------------------
+* vfat_valid_longname does not properly checked reserved names.
+* When a volume name is the same as a directory name in the root
+  directory of the filesystem, the directory name sometimes shows
+  up as an empty file.
+* autoconv option does not work correctly.
+
+BUG REPORTS
+----------------------------------------------------------------------
+If you have trouble with the VFAT filesystem, mail bug reports to
+chaffee@bmrc.cs.berkeley.edu.  Please specify the filename
+and the operation that gave you trouble.
+
+TEST SUITE
+----------------------------------------------------------------------
+If you plan to make any modifications to the vfat filesystem, please
+get the test suite that comes with the vfat distribution at
+
+  http://web.archive.org/web/*/http://bmrc.berkeley.edu/
+  people/chaffee/vfat.html
+
+This tests quite a few parts of the vfat filesystem and additional
+tests for new features or untested features would be appreciated.
+
+NOTES ON THE STRUCTURE OF THE VFAT FILESYSTEM
+----------------------------------------------------------------------
+(This documentation was provided by Galen C. Hunt <gchunt@cs.rochester.edu>
+ and lightly annotated by Gordon Chaffee).
+
+This document presents a very rough, technical overview of my
+knowledge of the extended FAT file system used in Windows NT 3.5 and
+Windows 95.  I don't guarantee that any of the following is correct,
+but it appears to be so.
+
+The extended FAT file system is almost identical to the FAT
+file system used in DOS versions up to and including 6.223410239847
+:-).  The significant change has been the addition of long file names.
+These names support up to 255 characters including spaces and lower
+case characters as opposed to the traditional 8.3 short names.
+
+Here is the description of the traditional FAT entry in the current
+Windows 95 filesystem:
+
+        struct directory { // Short 8.3 names 
+                unsigned char name[8];          // file name 
+                unsigned char ext[3];           // file extension 
+                unsigned char attr;             // attribute byte 
+		unsigned char lcase;		// Case for base and extension
+		unsigned char ctime_ms;		// Creation time, milliseconds
+		unsigned char ctime[2];		// Creation time
+		unsigned char cdate[2];		// Creation date
+		unsigned char adate[2];		// Last access date
+		unsigned char reserved[2];	// reserved values (ignored) 
+                unsigned char time[2];          // time stamp 
+                unsigned char date[2];          // date stamp 
+                unsigned char start[2];         // starting cluster number 
+                unsigned char size[4];          // size of the file 
+        };
+
+The lcase field specifies if the base and/or the extension of an 8.3
+name should be capitalized.  This field does not seem to be used by
+Windows 95 but it is used by Windows NT.  The case of filenames is not
+completely compatible from Windows NT to Windows 95.  It is not completely
+compatible in the reverse direction, however.  Filenames that fit in
+the 8.3 namespace and are written on Windows NT to be lowercase will
+show up as uppercase on Windows 95.
+
+Note that the "start" and "size" values are actually little
+endian integer values.  The descriptions of the fields in this
+structure are public knowledge and can be found elsewhere.
+
+With the extended FAT system, Microsoft has inserted extra
+directory entries for any files with extended names.  (Any name which
+legally fits within the old 8.3 encoding scheme does not have extra
+entries.)  I call these extra entries slots.  Basically, a slot is a
+specially formatted directory entry which holds up to 13 characters of
+a file's extended name.  Think of slots as additional labeling for the
+directory entry of the file to which they correspond.  Microsoft
+prefers to refer to the 8.3 entry for a file as its alias and the
+extended slot directory entries as the file name. 
+
+The C structure for a slot directory entry follows:
+
+        struct slot { // Up to 13 characters of a long name 
+                unsigned char id;               // sequence number for slot 
+                unsigned char name0_4[10];      // first 5 characters in name 
+                unsigned char attr;             // attribute byte
+                unsigned char reserved;         // always 0 
+                unsigned char alias_checksum;   // checksum for 8.3 alias 
+                unsigned char name5_10[12];     // 6 more characters in name
+                unsigned char start[2];         // starting cluster number
+                unsigned char name11_12[4];     // last 2 characters in name
+        };
+
+If the layout of the slots looks a little odd, it's only
+because of Microsoft's efforts to maintain compatibility with old
+software.  The slots must be disguised to prevent old software from
+panicking.  To this end, a number of measures are taken:
+
+        1) The attribute byte for a slot directory entry is always set
+           to 0x0f.  This corresponds to an old directory entry with
+           attributes of "hidden", "system", "read-only", and "volume
+           label".  Most old software will ignore any directory
+           entries with the "volume label" bit set.  Real volume label
+           entries don't have the other three bits set.
+
+        2) The starting cluster is always set to 0, an impossible
+           value for a DOS file.
+
+Because the extended FAT system is backward compatible, it is
+possible for old software to modify directory entries.  Measures must
+be taken to ensure the validity of slots.  An extended FAT system can
+verify that a slot does in fact belong to an 8.3 directory entry by
+the following:
+
+        1) Positioning.  Slots for a file always immediately proceed
+           their corresponding 8.3 directory entry.  In addition, each
+           slot has an id which marks its order in the extended file
+           name.  Here is a very abbreviated view of an 8.3 directory
+           entry and its corresponding long name slots for the file
+           "My Big File.Extension which is long":
+
+                <proceeding files...>
+                <slot #3, id = 0x43, characters = "h is long">
+                <slot #2, id = 0x02, characters = "xtension whic">
+                <slot #1, id = 0x01, characters = "My Big File.E">
+                <directory entry, name = "MYBIGFIL.EXT">
+
+           Note that the slots are stored from last to first.  Slots
+           are numbered from 1 to N.  The Nth slot is or'ed with 0x40
+           to mark it as the last one.
+
+        2) Checksum.  Each slot has an "alias_checksum" value.  The
+           checksum is calculated from the 8.3 name using the
+           following algorithm:
+
+                for (sum = i = 0; i < 11; i++) {
+                        sum = (((sum&1)<<7)|((sum&0xfe)>>1)) + name[i]
+                }
+
+	3) If there is free space in the final slot, a Unicode NULL (0x0000) 
+	   is stored after the final character.  After that, all unused 
+	   characters in the final slot are set to Unicode 0xFFFF.
+
+Finally, note that the extended name is stored in Unicode.  Each Unicode
+character takes either two or four bytes, UTF-16LE encoded.
diff --git a/Documentation/filesystems/vfs.txt b/Documentation/filesystems/vfs.txt
new file mode 100644
index 000000000..a6c6a8af4
--- /dev/null
+++ b/Documentation/filesystems/vfs.txt
@@ -0,0 +1,1253 @@
+
+	      Overview of the Linux Virtual File System
+
+	Original author: Richard Gooch <rgooch@atnf.csiro.au>
+
+		  Last updated on June 24, 2007.
+
+  Copyright (C) 1999 Richard Gooch
+  Copyright (C) 2005 Pekka Enberg
+
+  This file is released under the GPLv2.
+
+
+Introduction
+============
+
+The Virtual File System (also known as the Virtual Filesystem Switch)
+is the software layer in the kernel that provides the filesystem
+interface to userspace programs. It also provides an abstraction
+within the kernel which allows different filesystem implementations to
+coexist.
+
+VFS system calls open(2), stat(2), read(2), write(2), chmod(2) and so
+on are called from a process context. Filesystem locking is described
+in the document Documentation/filesystems/Locking.
+
+
+Directory Entry Cache (dcache)
+------------------------------
+
+The VFS implements the open(2), stat(2), chmod(2), and similar system
+calls. The pathname argument that is passed to them is used by the VFS
+to search through the directory entry cache (also known as the dentry
+cache or dcache). This provides a very fast look-up mechanism to
+translate a pathname (filename) into a specific dentry. Dentries live
+in RAM and are never saved to disc: they exist only for performance.
+
+The dentry cache is meant to be a view into your entire filespace. As
+most computers cannot fit all dentries in the RAM at the same time,
+some bits of the cache are missing. In order to resolve your pathname
+into a dentry, the VFS may have to resort to creating dentries along
+the way, and then loading the inode. This is done by looking up the
+inode.
+
+
+The Inode Object
+----------------
+
+An individual dentry usually has a pointer to an inode. Inodes are
+filesystem objects such as regular files, directories, FIFOs and other
+beasts.  They live either on the disc (for block device filesystems)
+or in the memory (for pseudo filesystems). Inodes that live on the
+disc are copied into the memory when required and changes to the inode
+are written back to disc. A single inode can be pointed to by multiple
+dentries (hard links, for example, do this).
+
+To look up an inode requires that the VFS calls the lookup() method of
+the parent directory inode. This method is installed by the specific
+filesystem implementation that the inode lives in. Once the VFS has
+the required dentry (and hence the inode), we can do all those boring
+things like open(2) the file, or stat(2) it to peek at the inode
+data. The stat(2) operation is fairly simple: once the VFS has the
+dentry, it peeks at the inode data and passes some of it back to
+userspace.
+
+
+The File Object
+---------------
+
+Opening a file requires another operation: allocation of a file
+structure (this is the kernel-side implementation of file
+descriptors). The freshly allocated file structure is initialized with
+a pointer to the dentry and a set of file operation member functions.
+These are taken from the inode data. The open() file method is then
+called so the specific filesystem implementation can do its work. You
+can see that this is another switch performed by the VFS. The file
+structure is placed into the file descriptor table for the process.
+
+Reading, writing and closing files (and other assorted VFS operations)
+is done by using the userspace file descriptor to grab the appropriate
+file structure, and then calling the required file structure method to
+do whatever is required. For as long as the file is open, it keeps the
+dentry in use, which in turn means that the VFS inode is still in use.
+
+
+Registering and Mounting a Filesystem
+=====================================
+
+To register and unregister a filesystem, use the following API
+functions:
+
+   #include <linux/fs.h>
+
+   extern int register_filesystem(struct file_system_type *);
+   extern int unregister_filesystem(struct file_system_type *);
+
+The passed struct file_system_type describes your filesystem. When a
+request is made to mount a filesystem onto a directory in your namespace,
+the VFS will call the appropriate mount() method for the specific
+filesystem.  New vfsmount referring to the tree returned by ->mount()
+will be attached to the mountpoint, so that when pathname resolution
+reaches the mountpoint it will jump into the root of that vfsmount.
+
+You can see all filesystems that are registered to the kernel in the
+file /proc/filesystems.
+
+
+struct file_system_type
+-----------------------
+
+This describes the filesystem. As of kernel 2.6.39, the following
+members are defined:
+
+struct file_system_type {
+	const char *name;
+	int fs_flags;
+        struct dentry *(*mount) (struct file_system_type *, int,
+                       const char *, void *);
+        void (*kill_sb) (struct super_block *);
+        struct module *owner;
+        struct file_system_type * next;
+        struct list_head fs_supers;
+	struct lock_class_key s_lock_key;
+	struct lock_class_key s_umount_key;
+};
+
+  name: the name of the filesystem type, such as "ext2", "iso9660",
+	"msdos" and so on
+
+  fs_flags: various flags (i.e. FS_REQUIRES_DEV, FS_NO_DCACHE, etc.)
+
+  mount: the method to call when a new instance of this
+	filesystem should be mounted
+
+  kill_sb: the method to call when an instance of this filesystem
+	should be shut down
+
+  owner: for internal VFS use: you should initialize this to THIS_MODULE in
+  	most cases.
+
+  next: for internal VFS use: you should initialize this to NULL
+
+  s_lock_key, s_umount_key: lockdep-specific
+
+The mount() method has the following arguments:
+
+  struct file_system_type *fs_type: describes the filesystem, partly initialized
+  	by the specific filesystem code
+
+  int flags: mount flags
+
+  const char *dev_name: the device name we are mounting.
+
+  void *data: arbitrary mount options, usually comes as an ASCII
+	string (see "Mount Options" section)
+
+The mount() method must return the root dentry of the tree requested by
+caller.  An active reference to its superblock must be grabbed and the
+superblock must be locked.  On failure it should return ERR_PTR(error).
+
+The arguments match those of mount(2) and their interpretation
+depends on filesystem type.  E.g. for block filesystems, dev_name is
+interpreted as block device name, that device is opened and if it
+contains a suitable filesystem image the method creates and initializes
+struct super_block accordingly, returning its root dentry to caller.
+
+->mount() may choose to return a subtree of existing filesystem - it
+doesn't have to create a new one.  The main result from the caller's
+point of view is a reference to dentry at the root of (sub)tree to
+be attached; creation of new superblock is a common side effect.
+
+The most interesting member of the superblock structure that the
+mount() method fills in is the "s_op" field. This is a pointer to
+a "struct super_operations" which describes the next level of the
+filesystem implementation.
+
+Usually, a filesystem uses one of the generic mount() implementations
+and provides a fill_super() callback instead. The generic variants are:
+
+  mount_bdev: mount a filesystem residing on a block device
+
+  mount_nodev: mount a filesystem that is not backed by a device
+
+  mount_single: mount a filesystem which shares the instance between
+  	all mounts
+
+A fill_super() callback implementation has the following arguments:
+
+  struct super_block *sb: the superblock structure. The callback
+  	must initialize this properly.
+
+  void *data: arbitrary mount options, usually comes as an ASCII
+	string (see "Mount Options" section)
+
+  int silent: whether or not to be silent on error
+
+
+The Superblock Object
+=====================
+
+A superblock object represents a mounted filesystem.
+
+
+struct super_operations
+-----------------------
+
+This describes how the VFS can manipulate the superblock of your
+filesystem. As of kernel 2.6.22, the following members are defined:
+
+struct super_operations {
+        struct inode *(*alloc_inode)(struct super_block *sb);
+        void (*destroy_inode)(struct inode *);
+
+        void (*dirty_inode) (struct inode *, int flags);
+        int (*write_inode) (struct inode *, int);
+        void (*drop_inode) (struct inode *);
+        void (*delete_inode) (struct inode *);
+        void (*put_super) (struct super_block *);
+        int (*sync_fs)(struct super_block *sb, int wait);
+        int (*freeze_fs) (struct super_block *);
+        int (*unfreeze_fs) (struct super_block *);
+        int (*statfs) (struct dentry *, struct kstatfs *);
+        int (*remount_fs) (struct super_block *, int *, char *);
+        void (*clear_inode) (struct inode *);
+        void (*umount_begin) (struct super_block *);
+
+        int (*show_options)(struct seq_file *, struct dentry *);
+
+        ssize_t (*quota_read)(struct super_block *, int, char *, size_t, loff_t);
+        ssize_t (*quota_write)(struct super_block *, int, const char *, size_t, loff_t);
+	int (*nr_cached_objects)(struct super_block *);
+	void (*free_cached_objects)(struct super_block *, int);
+};
+
+All methods are called without any locks being held, unless otherwise
+noted. This means that most methods can block safely. All methods are
+only called from a process context (i.e. not from an interrupt handler
+or bottom half).
+
+  alloc_inode: this method is called by alloc_inode() to allocate memory
+ 	for struct inode and initialize it.  If this function is not
+ 	defined, a simple 'struct inode' is allocated.  Normally
+ 	alloc_inode will be used to allocate a larger structure which
+ 	contains a 'struct inode' embedded within it.
+
+  destroy_inode: this method is called by destroy_inode() to release
+  	resources allocated for struct inode.  It is only required if
+  	->alloc_inode was defined and simply undoes anything done by
+	->alloc_inode.
+
+  dirty_inode: this method is called by the VFS to mark an inode dirty.
+
+  write_inode: this method is called when the VFS needs to write an
+	inode to disc.  The second parameter indicates whether the write
+	should be synchronous or not, not all filesystems check this flag.
+
+  drop_inode: called when the last access to the inode is dropped,
+	with the inode->i_lock spinlock held.
+
+	This method should be either NULL (normal UNIX filesystem
+	semantics) or "generic_delete_inode" (for filesystems that do not
+	want to cache inodes - causing "delete_inode" to always be
+	called regardless of the value of i_nlink)
+
+	The "generic_delete_inode()" behavior is equivalent to the
+	old practice of using "force_delete" in the put_inode() case,
+	but does not have the races that the "force_delete()" approach
+	had. 
+
+  delete_inode: called when the VFS wants to delete an inode
+
+  put_super: called when the VFS wishes to free the superblock
+	(i.e. unmount). This is called with the superblock lock held
+
+  sync_fs: called when VFS is writing out all dirty data associated with
+  	a superblock. The second parameter indicates whether the method
+	should wait until the write out has been completed. Optional.
+
+  freeze_fs: called when VFS is locking a filesystem and
+  	forcing it into a consistent state.  This method is currently
+  	used by the Logical Volume Manager (LVM).
+
+  unfreeze_fs: called when VFS is unlocking a filesystem and making it writable
+  	again.
+
+  statfs: called when the VFS needs to get filesystem statistics.
+
+  remount_fs: called when the filesystem is remounted. This is called
+	with the kernel lock held
+
+  clear_inode: called then the VFS clears the inode. Optional
+
+  umount_begin: called when the VFS is unmounting a filesystem.
+
+  show_options: called by the VFS to show mount options for
+	/proc/<pid>/mounts.  (see "Mount Options" section)
+
+  quota_read: called by the VFS to read from filesystem quota file.
+
+  quota_write: called by the VFS to write to filesystem quota file.
+
+  nr_cached_objects: called by the sb cache shrinking function for the
+	filesystem to return the number of freeable cached objects it contains.
+	Optional.
+
+  free_cache_objects: called by the sb cache shrinking function for the
+	filesystem to scan the number of objects indicated to try to free them.
+	Optional, but any filesystem implementing this method needs to also
+	implement ->nr_cached_objects for it to be called correctly.
+
+	We can't do anything with any errors that the filesystem might
+	encountered, hence the void return type. This will never be called if
+	the VM is trying to reclaim under GFP_NOFS conditions, hence this
+	method does not need to handle that situation itself.
+
+	Implementations must include conditional reschedule calls inside any
+	scanning loop that is done. This allows the VFS to determine
+	appropriate scan batch sizes without having to worry about whether
+	implementations will cause holdoff problems due to large scan batch
+	sizes.
+
+Whoever sets up the inode is responsible for filling in the "i_op" field. This
+is a pointer to a "struct inode_operations" which describes the methods that
+can be performed on individual inodes.
+
+struct xattr_handlers
+---------------------
+
+On filesystems that support extended attributes (xattrs), the s_xattr
+superblock field points to a NULL-terminated array of xattr handlers.  Extended
+attributes are name:value pairs.
+
+  name: Indicates that the handler matches attributes with the specified name
+	(such as "system.posix_acl_access"); the prefix field must be NULL.
+
+  prefix: Indicates that the handler matches all attributes with the specified
+	name prefix (such as "user."); the name field must be NULL.
+
+  list: Determine if attributes matching this xattr handler should be listed
+	for a particular dentry.  Used by some listxattr implementations like
+	generic_listxattr.
+
+  get: Called by the VFS to get the value of a particular extended attribute.
+	This method is called by the getxattr(2) system call.
+
+  set: Called by the VFS to set the value of a particular extended attribute.
+	When the new value is NULL, called to remove a particular extended
+	attribute.  This method is called by the the setxattr(2) and
+	removexattr(2) system calls.
+
+When none of the xattr handlers of a filesystem match the specified attribute
+name or when a filesystem doesn't support extended attributes, the various
+*xattr(2) system calls return -EOPNOTSUPP.
+
+
+The Inode Object
+================
+
+An inode object represents an object within the filesystem.
+
+
+struct inode_operations
+-----------------------
+
+This describes how the VFS can manipulate an inode in your
+filesystem. As of kernel 2.6.22, the following members are defined:
+
+struct inode_operations {
+	int (*create) (struct inode *,struct dentry *, umode_t, bool);
+	struct dentry * (*lookup) (struct inode *,struct dentry *, unsigned int);
+	int (*link) (struct dentry *,struct inode *,struct dentry *);
+	int (*unlink) (struct inode *,struct dentry *);
+	int (*symlink) (struct inode *,struct dentry *,const char *);
+	int (*mkdir) (struct inode *,struct dentry *,umode_t);
+	int (*rmdir) (struct inode *,struct dentry *);
+	int (*mknod) (struct inode *,struct dentry *,umode_t,dev_t);
+	int (*rename) (struct inode *, struct dentry *,
+			struct inode *, struct dentry *, unsigned int);
+	int (*readlink) (struct dentry *, char __user *,int);
+	const char *(*get_link) (struct dentry *, struct inode *,
+				 struct delayed_call *);
+	int (*permission) (struct inode *, int);
+	int (*get_acl)(struct inode *, int);
+	int (*setattr) (struct dentry *, struct iattr *);
+	int (*getattr) (const struct path *, struct kstat *, u32, unsigned int);
+	ssize_t (*listxattr) (struct dentry *, char *, size_t);
+	void (*update_time)(struct inode *, struct timespec *, int);
+	int (*atomic_open)(struct inode *, struct dentry *, struct file *,
+			unsigned open_flag, umode_t create_mode);
+	int (*tmpfile) (struct inode *, struct dentry *, umode_t);
+};
+
+Again, all methods are called without any locks being held, unless
+otherwise noted.
+
+  create: called by the open(2) and creat(2) system calls. Only
+	required if you want to support regular files. The dentry you
+	get should not have an inode (i.e. it should be a negative
+	dentry). Here you will probably call d_instantiate() with the
+	dentry and the newly created inode
+
+  lookup: called when the VFS needs to look up an inode in a parent
+	directory. The name to look for is found in the dentry. This
+	method must call d_add() to insert the found inode into the
+	dentry. The "i_count" field in the inode structure should be
+	incremented. If the named inode does not exist a NULL inode
+	should be inserted into the dentry (this is called a negative
+	dentry). Returning an error code from this routine must only
+	be done on a real error, otherwise creating inodes with system
+	calls like create(2), mknod(2), mkdir(2) and so on will fail.
+	If you wish to overload the dentry methods then you should
+	initialise the "d_dop" field in the dentry; this is a pointer
+	to a struct "dentry_operations".
+	This method is called with the directory inode semaphore held
+
+  link: called by the link(2) system call. Only required if you want
+	to support hard links. You will probably need to call
+	d_instantiate() just as you would in the create() method
+
+  unlink: called by the unlink(2) system call. Only required if you
+	want to support deleting inodes
+
+  symlink: called by the symlink(2) system call. Only required if you
+	want to support symlinks. You will probably need to call
+	d_instantiate() just as you would in the create() method
+
+  mkdir: called by the mkdir(2) system call. Only required if you want
+	to support creating subdirectories. You will probably need to
+	call d_instantiate() just as you would in the create() method
+
+  rmdir: called by the rmdir(2) system call. Only required if you want
+	to support deleting subdirectories
+
+  mknod: called by the mknod(2) system call to create a device (char,
+	block) inode or a named pipe (FIFO) or socket. Only required
+	if you want to support creating these types of inodes. You
+	will probably need to call d_instantiate() just as you would
+	in the create() method
+
+  rename: called by the rename(2) system call to rename the object to
+	have the parent and name given by the second inode and dentry.
+
+	The filesystem must return -EINVAL for any unsupported or
+	unknown	flags.  Currently the following flags are implemented:
+	(1) RENAME_NOREPLACE: this flag indicates that if the target
+	of the rename exists the rename should fail with -EEXIST
+	instead of replacing the target.  The VFS already checks for
+	existence, so for local filesystems the RENAME_NOREPLACE
+	implementation is equivalent to plain rename.
+	(2) RENAME_EXCHANGE: exchange source and target.  Both must
+	exist; this is checked by the VFS.  Unlike plain rename,
+	source and target may be of different type.
+
+  get_link: called by the VFS to follow a symbolic link to the
+	inode it points to.  Only required if you want to support
+	symbolic links.  This method returns the symlink body
+	to traverse (and possibly resets the current position with
+	nd_jump_link()).  If the body won't go away until the inode
+	is gone, nothing else is needed; if it needs to be otherwise
+	pinned, arrange for its release by having get_link(..., ..., done)
+	do set_delayed_call(done, destructor, argument).
+	In that case destructor(argument) will be called once VFS is
+	done with the body you've returned.
+	May be called in RCU mode; that is indicated by NULL dentry
+	argument.  If request can't be handled without leaving RCU mode,
+	have it return ERR_PTR(-ECHILD).
+
+  readlink: this is now just an override for use by readlink(2) for the
+	cases when ->get_link uses nd_jump_link() or object is not in
+	fact a symlink.  Normally filesystems should only implement
+	->get_link for symlinks and readlink(2) will automatically use
+	that.
+
+  permission: called by the VFS to check for access rights on a POSIX-like
+  	filesystem.
+
+	May be called in rcu-walk mode (mask & MAY_NOT_BLOCK). If in rcu-walk
+        mode, the filesystem must check the permission without blocking or
+	storing to the inode.
+
+	If a situation is encountered that rcu-walk cannot handle, return
+	-ECHILD and it will be called again in ref-walk mode.
+
+  setattr: called by the VFS to set attributes for a file. This method
+  	is called by chmod(2) and related system calls.
+
+  getattr: called by the VFS to get attributes of a file. This method
+  	is called by stat(2) and related system calls.
+
+  listxattr: called by the VFS to list all extended attributes for a
+	given file. This method is called by the listxattr(2) system call.
+
+  update_time: called by the VFS to update a specific time or the i_version of
+  	an inode.  If this is not defined the VFS will update the inode itself
+  	and call mark_inode_dirty_sync.
+
+  atomic_open: called on the last component of an open.  Using this optional
+  	method the filesystem can look up, possibly create and open the file in
+	one atomic operation.  If it wants to leave actual opening to the
+	caller (e.g. if the file turned out to be a symlink, device, or just
+	something filesystem won't do atomic open for), it may signal this by
+	returning finish_no_open(file, dentry).  This method is only called if
+	the last component is negative or needs lookup.  Cached positive dentries
+	are still handled by f_op->open().  If the file was created,
+	FMODE_CREATED flag should be set in file->f_mode.  In case of O_EXCL
+	the method must only succeed if the file didn't exist and hence FMODE_CREATED
+	shall always be set on success.
+
+  tmpfile: called in the end of O_TMPFILE open().  Optional, equivalent to
+	atomically creating, opening and unlinking a file in given directory.
+
+The Address Space Object
+========================
+
+The address space object is used to group and manage pages in the page
+cache.  It can be used to keep track of the pages in a file (or
+anything else) and also track the mapping of sections of the file into
+process address spaces.
+
+There are a number of distinct yet related services that an
+address-space can provide.  These include communicating memory
+pressure, page lookup by address, and keeping track of pages tagged as
+Dirty or Writeback.
+
+The first can be used independently to the others.  The VM can try to
+either write dirty pages in order to clean them, or release clean
+pages in order to reuse them.  To do this it can call the ->writepage
+method on dirty pages, and ->releasepage on clean pages with
+PagePrivate set. Clean pages without PagePrivate and with no external
+references will be released without notice being given to the
+address_space.
+
+To achieve this functionality, pages need to be placed on an LRU with
+lru_cache_add and mark_page_active needs to be called whenever the
+page is used.
+
+Pages are normally kept in a radix tree index by ->index. This tree
+maintains information about the PG_Dirty and PG_Writeback status of
+each page, so that pages with either of these flags can be found
+quickly.
+
+The Dirty tag is primarily used by mpage_writepages - the default
+->writepages method.  It uses the tag to find dirty pages to call
+->writepage on.  If mpage_writepages is not used (i.e. the address
+provides its own ->writepages) , the PAGECACHE_TAG_DIRTY tag is
+almost unused.  write_inode_now and sync_inode do use it (through
+__sync_single_inode) to check if ->writepages has been successful in
+writing out the whole address_space.
+
+The Writeback tag is used by filemap*wait* and sync_page* functions,
+via filemap_fdatawait_range, to wait for all writeback to complete.
+
+An address_space handler may attach extra information to a page,
+typically using the 'private' field in the 'struct page'.  If such
+information is attached, the PG_Private flag should be set.  This will
+cause various VM routines to make extra calls into the address_space
+handler to deal with that data.
+
+An address space acts as an intermediate between storage and
+application.  Data is read into the address space a whole page at a
+time, and provided to the application either by copying of the page,
+or by memory-mapping the page.
+Data is written into the address space by the application, and then
+written-back to storage typically in whole pages, however the
+address_space has finer control of write sizes.
+
+The read process essentially only requires 'readpage'.  The write
+process is more complicated and uses write_begin/write_end or
+set_page_dirty to write data into the address_space, and writepage
+and writepages to writeback data to storage.
+
+Adding and removing pages to/from an address_space is protected by the
+inode's i_mutex.
+
+When data is written to a page, the PG_Dirty flag should be set.  It
+typically remains set until writepage asks for it to be written.  This
+should clear PG_Dirty and set PG_Writeback.  It can be actually
+written at any point after PG_Dirty is clear.  Once it is known to be
+safe, PG_Writeback is cleared.
+
+Writeback makes use of a writeback_control structure to direct the
+operations.  This gives the the writepage and writepages operations some
+information about the nature of and reason for the writeback request,
+and the constraints under which it is being done.  It is also used to
+return information back to the caller about the result of a writepage or
+writepages request.
+
+Handling errors during writeback
+--------------------------------
+Most applications that do buffered I/O will periodically call a file
+synchronization call (fsync, fdatasync, msync or sync_file_range) to
+ensure that data written has made it to the backing store.  When there
+is an error during writeback, they expect that error to be reported when
+a file sync request is made.  After an error has been reported on one
+request, subsequent requests on the same file descriptor should return
+0, unless further writeback errors have occurred since the previous file
+syncronization.
+
+Ideally, the kernel would report errors only on file descriptions on
+which writes were done that subsequently failed to be written back.  The
+generic pagecache infrastructure does not track the file descriptions
+that have dirtied each individual page however, so determining which
+file descriptors should get back an error is not possible.
+
+Instead, the generic writeback error tracking infrastructure in the
+kernel settles for reporting errors to fsync on all file descriptions
+that were open at the time that the error occurred.  In a situation with
+multiple writers, all of them will get back an error on a subsequent fsync,
+even if all of the writes done through that particular file descriptor
+succeeded (or even if there were no writes on that file descriptor at all).
+
+Filesystems that wish to use this infrastructure should call
+mapping_set_error to record the error in the address_space when it
+occurs.  Then, after writing back data from the pagecache in their
+file->fsync operation, they should call file_check_and_advance_wb_err to
+ensure that the struct file's error cursor has advanced to the correct
+point in the stream of errors emitted by the backing device(s).
+
+struct address_space_operations
+-------------------------------
+
+This describes how the VFS can manipulate mapping of a file to page cache in
+your filesystem. The following members are defined:
+
+struct address_space_operations {
+	int (*writepage)(struct page *page, struct writeback_control *wbc);
+	int (*readpage)(struct file *, struct page *);
+	int (*writepages)(struct address_space *, struct writeback_control *);
+	int (*set_page_dirty)(struct page *page);
+	int (*readpages)(struct file *filp, struct address_space *mapping,
+			struct list_head *pages, unsigned nr_pages);
+	int (*write_begin)(struct file *, struct address_space *mapping,
+				loff_t pos, unsigned len, unsigned flags,
+				struct page **pagep, void **fsdata);
+	int (*write_end)(struct file *, struct address_space *mapping,
+				loff_t pos, unsigned len, unsigned copied,
+				struct page *page, void *fsdata);
+	sector_t (*bmap)(struct address_space *, sector_t);
+	void (*invalidatepage) (struct page *, unsigned int, unsigned int);
+	int (*releasepage) (struct page *, int);
+	void (*freepage)(struct page *);
+	ssize_t (*direct_IO)(struct kiocb *, struct iov_iter *iter);
+	/* isolate a page for migration */
+	bool (*isolate_page) (struct page *, isolate_mode_t);
+	/* migrate the contents of a page to the specified target */
+	int (*migratepage) (struct page *, struct page *);
+	/* put migration-failed page back to right list */
+	void (*putback_page) (struct page *);
+	int (*launder_page) (struct page *);
+
+	int (*is_partially_uptodate) (struct page *, unsigned long,
+					unsigned long);
+	void (*is_dirty_writeback) (struct page *, bool *, bool *);
+	int (*error_remove_page) (struct mapping *mapping, struct page *page);
+	int (*swap_activate)(struct file *);
+	int (*swap_deactivate)(struct file *);
+};
+
+  writepage: called by the VM to write a dirty page to backing store.
+      This may happen for data integrity reasons (i.e. 'sync'), or
+      to free up memory (flush).  The difference can be seen in
+      wbc->sync_mode.
+      The PG_Dirty flag has been cleared and PageLocked is true.
+      writepage should start writeout, should set PG_Writeback,
+      and should make sure the page is unlocked, either synchronously
+      or asynchronously when the write operation completes.
+
+      If wbc->sync_mode is WB_SYNC_NONE, ->writepage doesn't have to
+      try too hard if there are problems, and may choose to write out
+      other pages from the mapping if that is easier (e.g. due to
+      internal dependencies).  If it chooses not to start writeout, it
+      should return AOP_WRITEPAGE_ACTIVATE so that the VM will not keep
+      calling ->writepage on that page.
+
+      See the file "Locking" for more details.
+
+  readpage: called by the VM to read a page from backing store.
+       The page will be Locked when readpage is called, and should be
+       unlocked and marked uptodate once the read completes.
+       If ->readpage discovers that it needs to unlock the page for
+       some reason, it can do so, and then return AOP_TRUNCATED_PAGE.
+       In this case, the page will be relocated, relocked and if
+       that all succeeds, ->readpage will be called again.
+
+  writepages: called by the VM to write out pages associated with the
+  	address_space object.  If wbc->sync_mode is WBC_SYNC_ALL, then
+  	the writeback_control will specify a range of pages that must be
+  	written out.  If it is WBC_SYNC_NONE, then a nr_to_write is given
+	and that many pages should be written if possible.
+	If no ->writepages is given, then mpage_writepages is used
+  	instead.  This will choose pages from the address space that are
+  	tagged as DIRTY and will pass them to ->writepage.
+
+  set_page_dirty: called by the VM to set a page dirty.
+        This is particularly needed if an address space attaches
+        private data to a page, and that data needs to be updated when
+        a page is dirtied.  This is called, for example, when a memory
+	mapped page gets modified.
+	If defined, it should set the PageDirty flag, and the
+        PAGECACHE_TAG_DIRTY tag in the radix tree.
+
+  readpages: called by the VM to read pages associated with the address_space
+  	object. This is essentially just a vector version of
+  	readpage.  Instead of just one page, several pages are
+  	requested.
+	readpages is only used for read-ahead, so read errors are
+  	ignored.  If anything goes wrong, feel free to give up.
+
+  write_begin:
+	Called by the generic buffered write code to ask the filesystem to
+	prepare to write len bytes at the given offset in the file. The
+	address_space should check that the write will be able to complete,
+	by allocating space if necessary and doing any other internal
+	housekeeping.  If the write will update parts of any basic-blocks on
+	storage, then those blocks should be pre-read (if they haven't been
+	read already) so that the updated blocks can be written out properly.
+
+        The filesystem must return the locked pagecache page for the specified
+	offset, in *pagep, for the caller to write into.
+
+	It must be able to cope with short writes (where the length passed to
+	write_begin is greater than the number of bytes copied into the page).
+
+	flags is a field for AOP_FLAG_xxx flags, described in
+	include/linux/fs.h.
+
+        A void * may be returned in fsdata, which then gets passed into
+        write_end.
+
+        Returns 0 on success; < 0 on failure (which is the error code), in
+	which case write_end is not called.
+
+  write_end: After a successful write_begin, and data copy, write_end must
+        be called. len is the original len passed to write_begin, and copied
+        is the amount that was able to be copied.
+
+        The filesystem must take care of unlocking the page and releasing it
+        refcount, and updating i_size.
+
+        Returns < 0 on failure, otherwise the number of bytes (<= 'copied')
+        that were able to be copied into pagecache.
+
+  bmap: called by the VFS to map a logical block offset within object to
+  	physical block number. This method is used by the FIBMAP
+  	ioctl and for working with swap-files.  To be able to swap to
+  	a file, the file must have a stable mapping to a block
+  	device.  The swap system does not go through the filesystem
+  	but instead uses bmap to find out where the blocks in the file
+  	are and uses those addresses directly.
+
+  invalidatepage: If a page has PagePrivate set, then invalidatepage
+        will be called when part or all of the page is to be removed
+	from the address space.  This generally corresponds to either a
+	truncation, punch hole  or a complete invalidation of the address
+	space (in the latter case 'offset' will always be 0 and 'length'
+	will be PAGE_SIZE). Any private data associated with the page
+	should be updated to reflect this truncation.  If offset is 0 and
+	length is PAGE_SIZE, then the private data should be released,
+	because the page must be able to be completely discarded.  This may
+	be done by calling the ->releasepage function, but in this case the
+	release MUST succeed.
+
+  releasepage: releasepage is called on PagePrivate pages to indicate
+        that the page should be freed if possible.  ->releasepage
+        should remove any private data from the page and clear the
+        PagePrivate flag. If releasepage() fails for some reason, it must
+	indicate failure with a 0 return value.
+	releasepage() is used in two distinct though related cases.  The
+	first is when the VM finds a clean page with no active users and
+        wants to make it a free page.  If ->releasepage succeeds, the
+        page will be removed from the address_space and become free.
+
+	The second case is when a request has been made to invalidate
+        some or all pages in an address_space.  This can happen
+        through the fadvise(POSIX_FADV_DONTNEED) system call or by the
+        filesystem explicitly requesting it as nfs and 9fs do (when
+        they believe the cache may be out of date with storage) by
+        calling invalidate_inode_pages2().
+	If the filesystem makes such a call, and needs to be certain
+        that all pages are invalidated, then its releasepage will
+        need to ensure this.  Possibly it can clear the PageUptodate
+        bit if it cannot free private data yet.
+
+  freepage: freepage is called once the page is no longer visible in
+        the page cache in order to allow the cleanup of any private
+	data. Since it may be called by the memory reclaimer, it
+	should not assume that the original address_space mapping still
+	exists, and it should not block.
+
+  direct_IO: called by the generic read/write routines to perform
+        direct_IO - that is IO requests which bypass the page cache
+        and transfer data directly between the storage and the
+        application's address space.
+
+  isolate_page: Called by the VM when isolating a movable non-lru page.
+	If page is successfully isolated, VM marks the page as PG_isolated
+	via __SetPageIsolated.
+
+  migrate_page:  This is used to compact the physical memory usage.
+        If the VM wants to relocate a page (maybe off a memory card
+        that is signalling imminent failure) it will pass a new page
+	and an old page to this function.  migrate_page should
+	transfer any private data across and update any references
+        that it has to the page.
+
+  putback_page: Called by the VM when isolated page's migration fails.
+
+  launder_page: Called before freeing a page - it writes back the dirty page. To
+  	prevent redirtying the page, it is kept locked during the whole
+	operation.
+
+  is_partially_uptodate: Called by the VM when reading a file through the
+	pagecache when the underlying blocksize != pagesize. If the required
+	block is up to date then the read can complete without needing the IO
+	to bring the whole page up to date.
+
+  is_dirty_writeback: Called by the VM when attempting to reclaim a page.
+	The VM uses dirty and writeback information to determine if it needs
+	to stall to allow flushers a chance to complete some IO. Ordinarily
+	it can use PageDirty and PageWriteback but some filesystems have
+	more complex state (unstable pages in NFS prevent reclaim) or
+	do not set those flags due to locking problems. This callback
+	allows a filesystem to indicate to the VM if a page should be
+	treated as dirty or writeback for the purposes of stalling.
+
+  error_remove_page: normally set to generic_error_remove_page if truncation
+	is ok for this address space. Used for memory failure handling.
+	Setting this implies you deal with pages going away under you,
+	unless you have them locked or reference counts increased.
+
+  swap_activate: Called when swapon is used on a file to allocate
+	space if necessary and pin the block lookup information in
+	memory. A return value of zero indicates success,
+	in which case this file can be used to back swapspace.
+
+  swap_deactivate: Called during swapoff on files where swap_activate
+	was successful.
+
+
+The File Object
+===============
+
+A file object represents a file opened by a process. This is also known
+as an "open file description" in POSIX parlance.
+
+
+struct file_operations
+----------------------
+
+This describes how the VFS can manipulate an open file. As of kernel
+4.18, the following members are defined:
+
+struct file_operations {
+	struct module *owner;
+	loff_t (*llseek) (struct file *, loff_t, int);
+	ssize_t (*read) (struct file *, char __user *, size_t, loff_t *);
+	ssize_t (*write) (struct file *, const char __user *, size_t, loff_t *);
+	ssize_t (*read_iter) (struct kiocb *, struct iov_iter *);
+	ssize_t (*write_iter) (struct kiocb *, struct iov_iter *);
+	int (*iterate) (struct file *, struct dir_context *);
+	int (*iterate_shared) (struct file *, struct dir_context *);
+	__poll_t (*poll) (struct file *, struct poll_table_struct *);
+	long (*unlocked_ioctl) (struct file *, unsigned int, unsigned long);
+	long (*compat_ioctl) (struct file *, unsigned int, unsigned long);
+	int (*mmap) (struct file *, struct vm_area_struct *);
+	int (*open) (struct inode *, struct file *);
+	int (*flush) (struct file *, fl_owner_t id);
+	int (*release) (struct inode *, struct file *);
+	int (*fsync) (struct file *, loff_t, loff_t, int datasync);
+	int (*fasync) (int, struct file *, int);
+	int (*lock) (struct file *, int, struct file_lock *);
+	ssize_t (*sendpage) (struct file *, struct page *, int, size_t, loff_t *, int);
+	unsigned long (*get_unmapped_area)(struct file *, unsigned long, unsigned long, unsigned long, unsigned long);
+	int (*check_flags)(int);
+	int (*flock) (struct file *, int, struct file_lock *);
+	ssize_t (*splice_write)(struct pipe_inode_info *, struct file *, loff_t *, size_t, unsigned int);
+	ssize_t (*splice_read)(struct file *, loff_t *, struct pipe_inode_info *, size_t, unsigned int);
+	int (*setlease)(struct file *, long, struct file_lock **, void **);
+	long (*fallocate)(struct file *file, int mode, loff_t offset,
+			  loff_t len);
+	void (*show_fdinfo)(struct seq_file *m, struct file *f);
+#ifndef CONFIG_MMU
+	unsigned (*mmap_capabilities)(struct file *);
+#endif
+	ssize_t (*copy_file_range)(struct file *, loff_t, struct file *, loff_t, size_t, unsigned int);
+	int (*clone_file_range)(struct file *, loff_t, struct file *, loff_t, u64);
+	int (*dedupe_file_range)(struct file *, loff_t, struct file *, loff_t, u64);
+	int (*fadvise)(struct file *, loff_t, loff_t, int);
+};
+
+Again, all methods are called without any locks being held, unless
+otherwise noted.
+
+  llseek: called when the VFS needs to move the file position index
+
+  read: called by read(2) and related system calls
+
+  read_iter: possibly asynchronous read with iov_iter as destination
+
+  write: called by write(2) and related system calls
+
+  write_iter: possibly asynchronous write with iov_iter as source
+
+  iterate: called when the VFS needs to read the directory contents
+
+  iterate_shared: called when the VFS needs to read the directory contents
+	when filesystem supports concurrent dir iterators
+
+  poll: called by the VFS when a process wants to check if there is
+	activity on this file and (optionally) go to sleep until there
+	is activity. Called by the select(2) and poll(2) system calls
+
+  unlocked_ioctl: called by the ioctl(2) system call.
+
+  compat_ioctl: called by the ioctl(2) system call when 32 bit system calls
+ 	 are used on 64 bit kernels.
+
+  mmap: called by the mmap(2) system call
+
+  open: called by the VFS when an inode should be opened. When the VFS
+	opens a file, it creates a new "struct file". It then calls the
+	open method for the newly allocated file structure. You might
+	think that the open method really belongs in
+	"struct inode_operations", and you may be right. I think it's
+	done the way it is because it makes filesystems simpler to
+	implement. The open() method is a good place to initialize the
+	"private_data" member in the file structure if you want to point
+	to a device structure
+
+  flush: called by the close(2) system call to flush a file
+
+  release: called when the last reference to an open file is closed
+
+  fsync: called by the fsync(2) system call. Also see the section above
+	 entitled "Handling errors during writeback".
+
+  fasync: called by the fcntl(2) system call when asynchronous
+	(non-blocking) mode is enabled for a file
+
+  lock: called by the fcntl(2) system call for F_GETLK, F_SETLK, and F_SETLKW
+  	commands
+
+  get_unmapped_area: called by the mmap(2) system call
+
+  check_flags: called by the fcntl(2) system call for F_SETFL command
+
+  flock: called by the flock(2) system call
+
+  splice_write: called by the VFS to splice data from a pipe to a file. This
+		method is used by the splice(2) system call
+
+  splice_read: called by the VFS to splice data from file to a pipe. This
+	       method is used by the splice(2) system call
+
+  setlease: called by the VFS to set or release a file lock lease. setlease
+	    implementations should call generic_setlease to record or remove
+	    the lease in the inode after setting it.
+
+  fallocate: called by the VFS to preallocate blocks or punch a hole.
+
+  copy_file_range: called by the copy_file_range(2) system call.
+
+  clone_file_range: called by the ioctl(2) system call for FICLONERANGE and
+	FICLONE commands.
+
+  dedupe_file_range: called by the ioctl(2) system call for FIDEDUPERANGE
+	command.
+
+  fadvise: possibly called by the fadvise64() system call.
+
+Note that the file operations are implemented by the specific
+filesystem in which the inode resides. When opening a device node
+(character or block special) most filesystems will call special
+support routines in the VFS which will locate the required device
+driver information. These support routines replace the filesystem file
+operations with those for the device driver, and then proceed to call
+the new open() method for the file. This is how opening a device file
+in the filesystem eventually ends up calling the device driver open()
+method.
+
+
+Directory Entry Cache (dcache)
+==============================
+
+
+struct dentry_operations
+------------------------
+
+This describes how a filesystem can overload the standard dentry
+operations. Dentries and the dcache are the domain of the VFS and the
+individual filesystem implementations. Device drivers have no business
+here. These methods may be set to NULL, as they are either optional or
+the VFS uses a default. As of kernel 2.6.22, the following members are
+defined:
+
+struct dentry_operations {
+	int (*d_revalidate)(struct dentry *, unsigned int);
+	int (*d_weak_revalidate)(struct dentry *, unsigned int);
+	int (*d_hash)(const struct dentry *, struct qstr *);
+	int (*d_compare)(const struct dentry *,
+			unsigned int, const char *, const struct qstr *);
+	int (*d_delete)(const struct dentry *);
+	int (*d_init)(struct dentry *);
+	void (*d_release)(struct dentry *);
+	void (*d_iput)(struct dentry *, struct inode *);
+	char *(*d_dname)(struct dentry *, char *, int);
+	struct vfsmount *(*d_automount)(struct path *);
+	int (*d_manage)(const struct path *, bool);
+	struct dentry *(*d_real)(struct dentry *, const struct inode *);
+};
+
+  d_revalidate: called when the VFS needs to revalidate a dentry. This
+	is called whenever a name look-up finds a dentry in the
+	dcache. Most local filesystems leave this as NULL, because all their
+	dentries in the dcache are valid. Network filesystems are different
+	since things can change on the server without the client necessarily
+	being aware of it.
+
+	This function should return a positive value if the dentry is still
+	valid, and zero or a negative error code if it isn't.
+
+	d_revalidate may be called in rcu-walk mode (flags & LOOKUP_RCU).
+	If in rcu-walk mode, the filesystem must revalidate the dentry without
+	blocking or storing to the dentry, d_parent and d_inode should not be
+	used without care (because they can change and, in d_inode case, even
+	become NULL under us).
+
+	If a situation is encountered that rcu-walk cannot handle, return
+	-ECHILD and it will be called again in ref-walk mode.
+
+ d_weak_revalidate: called when the VFS needs to revalidate a "jumped" dentry.
+	This is called when a path-walk ends at dentry that was not acquired by
+	doing a lookup in the parent directory. This includes "/", "." and "..",
+	as well as procfs-style symlinks and mountpoint traversal.
+
+	In this case, we are less concerned with whether the dentry is still
+	fully correct, but rather that the inode is still valid. As with
+	d_revalidate, most local filesystems will set this to NULL since their
+	dcache entries are always valid.
+
+	This function has the same return code semantics as d_revalidate.
+
+	d_weak_revalidate is only called after leaving rcu-walk mode.
+
+  d_hash: called when the VFS adds a dentry to the hash table. The first
+	dentry passed to d_hash is the parent directory that the name is
+	to be hashed into.
+
+	Same locking and synchronisation rules as d_compare regarding
+	what is safe to dereference etc.
+
+  d_compare: called to compare a dentry name with a given name. The first
+	dentry is the parent of the dentry to be compared, the second is
+	the child dentry. len and name string are properties of the dentry
+	to be compared. qstr is the name to compare it with.
+
+	Must be constant and idempotent, and should not take locks if
+	possible, and should not or store into the dentry.
+	Should not dereference pointers outside the dentry without
+	lots of care (eg.  d_parent, d_inode, d_name should not be used).
+
+	However, our vfsmount is pinned, and RCU held, so the dentries and
+	inodes won't disappear, neither will our sb or filesystem module.
+	->d_sb may be used.
+
+	It is a tricky calling convention because it needs to be called under
+	"rcu-walk", ie. without any locks or references on things.
+
+  d_delete: called when the last reference to a dentry is dropped and the
+	dcache is deciding whether or not to cache it. Return 1 to delete
+	immediately, or 0 to cache the dentry. Default is NULL which means to
+	always cache a reachable dentry. d_delete must be constant and
+	idempotent.
+
+  d_init: called when a dentry is allocated
+
+  d_release: called when a dentry is really deallocated
+
+  d_iput: called when a dentry loses its inode (just prior to its
+	being deallocated). The default when this is NULL is that the
+	VFS calls iput(). If you define this method, you must call
+	iput() yourself
+
+  d_dname: called when the pathname of a dentry should be generated.
+	Useful for some pseudo filesystems (sockfs, pipefs, ...) to delay
+	pathname generation. (Instead of doing it when dentry is created,
+	it's done only when the path is needed.). Real filesystems probably
+	dont want to use it, because their dentries are present in global
+	dcache hash, so their hash should be an invariant. As no lock is
+	held, d_dname() should not try to modify the dentry itself, unless
+	appropriate SMP safety is used. CAUTION : d_path() logic is quite
+	tricky. The correct way to return for example "Hello" is to put it
+	at the end of the buffer, and returns a pointer to the first char.
+	dynamic_dname() helper function is provided to take care of this.
+
+	Example :
+
+	static char *pipefs_dname(struct dentry *dent, char *buffer, int buflen)
+	{
+		return dynamic_dname(dentry, buffer, buflen, "pipe:[%lu]",
+				dentry->d_inode->i_ino);
+	}
+
+  d_automount: called when an automount dentry is to be traversed (optional).
+	This should create a new VFS mount record and return the record to the
+	caller.  The caller is supplied with a path parameter giving the
+	automount directory to describe the automount target and the parent
+	VFS mount record to provide inheritable mount parameters.  NULL should
+	be returned if someone else managed to make the automount first.  If
+	the vfsmount creation failed, then an error code should be returned.
+	If -EISDIR is returned, then the directory will be treated as an
+	ordinary directory and returned to pathwalk to continue walking.
+
+	If a vfsmount is returned, the caller will attempt to mount it on the
+	mountpoint and will remove the vfsmount from its expiration list in
+	the case of failure.  The vfsmount should be returned with 2 refs on
+	it to prevent automatic expiration - the caller will clean up the
+	additional ref.
+
+	This function is only used if DCACHE_NEED_AUTOMOUNT is set on the
+	dentry.  This is set by __d_instantiate() if S_AUTOMOUNT is set on the
+	inode being added.
+
+  d_manage: called to allow the filesystem to manage the transition from a
+	dentry (optional).  This allows autofs, for example, to hold up clients
+	waiting to explore behind a 'mountpoint' whilst letting the daemon go
+	past and construct the subtree there.  0 should be returned to let the
+	calling process continue.  -EISDIR can be returned to tell pathwalk to
+	use this directory as an ordinary directory and to ignore anything
+	mounted on it and not to check the automount flag.  Any other error
+	code will abort pathwalk completely.
+
+	If the 'rcu_walk' parameter is true, then the caller is doing a
+	pathwalk in RCU-walk mode.  Sleeping is not permitted in this mode,
+	and the caller can be asked to leave it and call again by returning
+	-ECHILD.  -EISDIR may also be returned to tell pathwalk to
+	ignore d_automount or any mounts.
+
+	This function is only used if DCACHE_MANAGE_TRANSIT is set on the
+	dentry being transited from.
+
+  d_real: overlay/union type filesystems implement this method to return one of
+	the underlying dentries hidden by the overlay.  It is used in two
+	different modes:
+
+	Called from file_dentry() it returns the real dentry matching the inode
+	argument.  The real dentry may be from a lower layer already copied up,
+	but still referenced from the file.  This mode is selected with a
+	non-NULL inode argument.
+
+	With NULL inode the topmost real underlying dentry is returned.
+
+Each dentry has a pointer to its parent dentry, as well as a hash list
+of child dentries. Child dentries are basically like files in a
+directory.
+
+
+Directory Entry Cache API
+--------------------------
+
+There are a number of functions defined which permit a filesystem to
+manipulate dentries:
+
+  dget: open a new handle for an existing dentry (this just increments
+	the usage count)
+
+  dput: close a handle for a dentry (decrements the usage count). If
+	the usage count drops to 0, and the dentry is still in its
+	parent's hash, the "d_delete" method is called to check whether
+	it should be cached. If it should not be cached, or if the dentry
+	is not hashed, it is deleted. Otherwise cached dentries are put
+	into an LRU list to be reclaimed on memory shortage.
+
+  d_drop: this unhashes a dentry from its parents hash list. A
+	subsequent call to dput() will deallocate the dentry if its
+	usage count drops to 0
+
+  d_delete: delete a dentry. If there are no other open references to
+	the dentry then the dentry is turned into a negative dentry
+	(the d_iput() method is called). If there are other
+	references, then d_drop() is called instead
+
+  d_add: add a dentry to its parents hash list and then calls
+	d_instantiate()
+
+  d_instantiate: add a dentry to the alias hash list for the inode and
+	updates the "d_inode" member. The "i_count" member in the
+	inode structure should be set/incremented. If the inode
+	pointer is NULL, the dentry is called a "negative
+	dentry". This function is commonly called when an inode is
+	created for an existing negative dentry
+
+  d_lookup: look up a dentry given its parent and path name component
+	It looks up the child of that given name from the dcache
+	hash table. If it is found, the reference count is incremented
+	and the dentry is returned. The caller must use dput()
+	to free the dentry when it finishes using it.
+
+Mount Options
+=============
+
+Parsing options
+---------------
+
+On mount and remount the filesystem is passed a string containing a
+comma separated list of mount options.  The options can have either of
+these forms:
+
+  option
+  option=value
+
+The <linux/parser.h> header defines an API that helps parse these
+options.  There are plenty of examples on how to use it in existing
+filesystems.
+
+Showing options
+---------------
+
+If a filesystem accepts mount options, it must define show_options()
+to show all the currently active options.  The rules are:
+
+  - options MUST be shown which are not default or their values differ
+    from the default
+
+  - options MAY be shown which are enabled by default or have their
+    default value
+
+Options used only internally between a mount helper and the kernel
+(such as file descriptors), or which only have an effect during the
+mounting (such as ones controlling the creation of a journal) are exempt
+from the above rules.
+
+The underlying reason for the above rules is to make sure, that a
+mount can be accurately replicated (e.g. umounting and mounting again)
+based on the information found in /proc/mounts.
+
+Resources
+=========
+
+(Note some of these resources are not up-to-date with the latest kernel
+ version.)
+
+Creating Linux virtual filesystems. 2002
+    <http://lwn.net/Articles/13325/>
+
+The Linux Virtual File-system Layer by Neil Brown. 1999
+    <http://www.cse.unsw.edu.au/~neilb/oss/linux-commentary/vfs.html>
+
+A tour of the Linux VFS by Michael K. Johnson. 1996
+    <http://www.tldp.org/LDP/khg/HyperNews/get/fs/vfstour.html>
+
+A small trail through the Linux kernel by Andries Brouwer. 2001
+    <http://www.win.tue.nl/~aeb/linux/vfs/trail.html>
diff --git a/Documentation/filesystems/xfs-delayed-logging-design.txt b/Documentation/filesystems/xfs-delayed-logging-design.txt
new file mode 100644
index 000000000..2ce36439c
--- /dev/null
+++ b/Documentation/filesystems/xfs-delayed-logging-design.txt
@@ -0,0 +1,793 @@
+XFS Delayed Logging Design
+--------------------------
+
+Introduction to Re-logging in XFS
+---------------------------------
+
+XFS logging is a combination of logical and physical logging. Some objects,
+such as inodes and dquots, are logged in logical format where the details
+logged are made up of the changes to in-core structures rather than on-disk
+structures. Other objects - typically buffers - have their physical changes
+logged. The reason for these differences is to reduce the amount of log space
+required for objects that are frequently logged. Some parts of inodes are more
+frequently logged than others, and inodes are typically more frequently logged
+than any other object (except maybe the superblock buffer) so keeping the
+amount of metadata logged low is of prime importance.
+
+The reason that this is such a concern is that XFS allows multiple separate
+modifications to a single object to be carried in the log at any given time.
+This allows the log to avoid needing to flush each change to disk before
+recording a new change to the object. XFS does this via a method called
+"re-logging". Conceptually, this is quite simple - all it requires is that any
+new change to the object is recorded with a *new copy* of all the existing
+changes in the new transaction that is written to the log.
+
+That is, if we have a sequence of changes A through to F, and the object was
+written to disk after change D, we would see in the log the following series
+of transactions, their contents and the log sequence number (LSN) of the
+transaction:
+
+	Transaction		Contents	LSN
+	   A			   A		   X
+	   B			  A+B		  X+n
+	   C			 A+B+C		 X+n+m
+	   D			A+B+C+D		X+n+m+o
+	    <object written to disk>
+	   E			   E		   Y (> X+n+m+o)
+	   F			  E+F		  Yٍ+p
+
+In other words, each time an object is relogged, the new transaction contains
+the aggregation of all the previous changes currently held only in the log.
+
+This relogging technique also allows objects to be moved forward in the log so
+that an object being relogged does not prevent the tail of the log from ever
+moving forward.  This can be seen in the table above by the changing
+(increasing) LSN of each subsequent transaction - the LSN is effectively a
+direct encoding of the location in the log of the transaction.
+
+This relogging is also used to implement long-running, multiple-commit
+transactions.  These transaction are known as rolling transactions, and require
+a special log reservation known as a permanent transaction reservation. A
+typical example of a rolling transaction is the removal of extents from an
+inode which can only be done at a rate of two extents per transaction because
+of reservation size limitations. Hence a rolling extent removal transaction
+keeps relogging the inode and btree buffers as they get modified in each
+removal operation. This keeps them moving forward in the log as the operation
+progresses, ensuring that current operation never gets blocked by itself if the
+log wraps around.
+
+Hence it can be seen that the relogging operation is fundamental to the correct
+working of the XFS journalling subsystem. From the above description, most
+people should be able to see why the XFS metadata operations writes so much to
+the log - repeated operations to the same objects write the same changes to
+the log over and over again. Worse is the fact that objects tend to get
+dirtier as they get relogged, so each subsequent transaction is writing more
+metadata into the log.
+
+Another feature of the XFS transaction subsystem is that most transactions are
+asynchronous. That is, they don't commit to disk until either a log buffer is
+filled (a log buffer can hold multiple transactions) or a synchronous operation
+forces the log buffers holding the transactions to disk. This means that XFS is
+doing aggregation of transactions in memory - batching them, if you like - to
+minimise the impact of the log IO on transaction throughput.
+
+The limitation on asynchronous transaction throughput is the number and size of
+log buffers made available by the log manager. By default there are 8 log
+buffers available and the size of each is 32kB - the size can be increased up
+to 256kB by use of a mount option.
+
+Effectively, this gives us the maximum bound of outstanding metadata changes
+that can be made to the filesystem at any point in time - if all the log
+buffers are full and under IO, then no more transactions can be committed until
+the current batch completes. It is now common for a single current CPU core to
+be to able to issue enough transactions to keep the log buffers full and under
+IO permanently. Hence the XFS journalling subsystem can be considered to be IO
+bound.
+
+Delayed Logging: Concepts
+-------------------------
+
+The key thing to note about the asynchronous logging combined with the
+relogging technique XFS uses is that we can be relogging changed objects
+multiple times before they are committed to disk in the log buffers. If we
+return to the previous relogging example, it is entirely possible that
+transactions A through D are committed to disk in the same log buffer.
+
+That is, a single log buffer may contain multiple copies of the same object,
+but only one of those copies needs to be there - the last one "D", as it
+contains all the changes from the previous changes. In other words, we have one
+necessary copy in the log buffer, and three stale copies that are simply
+wasting space. When we are doing repeated operations on the same set of
+objects, these "stale objects" can be over 90% of the space used in the log
+buffers. It is clear that reducing the number of stale objects written to the
+log would greatly reduce the amount of metadata we write to the log, and this
+is the fundamental goal of delayed logging.
+
+From a conceptual point of view, XFS is already doing relogging in memory (where
+memory == log buffer), only it is doing it extremely inefficiently. It is using
+logical to physical formatting to do the relogging because there is no
+infrastructure to keep track of logical changes in memory prior to physically
+formatting the changes in a transaction to the log buffer. Hence we cannot avoid
+accumulating stale objects in the log buffers.
+
+Delayed logging is the name we've given to keeping and tracking transactional
+changes to objects in memory outside the log buffer infrastructure. Because of
+the relogging concept fundamental to the XFS journalling subsystem, this is
+actually relatively easy to do - all the changes to logged items are already
+tracked in the current infrastructure. The big problem is how to accumulate
+them and get them to the log in a consistent, recoverable manner.
+Describing the problems and how they have been solved is the focus of this
+document.
+
+One of the key changes that delayed logging makes to the operation of the
+journalling subsystem is that it disassociates the amount of outstanding
+metadata changes from the size and number of log buffers available. In other
+words, instead of there only being a maximum of 2MB of transaction changes not
+written to the log at any point in time, there may be a much greater amount
+being accumulated in memory. Hence the potential for loss of metadata on a
+crash is much greater than for the existing logging mechanism.
+
+It should be noted that this does not change the guarantee that log recovery
+will result in a consistent filesystem. What it does mean is that as far as the
+recovered filesystem is concerned, there may be many thousands of transactions
+that simply did not occur as a result of the crash. This makes it even more
+important that applications that care about their data use fsync() where they
+need to ensure application level data integrity is maintained.
+
+It should be noted that delayed logging is not an innovative new concept that
+warrants rigorous proofs to determine whether it is correct or not. The method
+of accumulating changes in memory for some period before writing them to the
+log is used effectively in many filesystems including ext3 and ext4. Hence
+no time is spent in this document trying to convince the reader that the
+concept is sound. Instead it is simply considered a "solved problem" and as
+such implementing it in XFS is purely an exercise in software engineering.
+
+The fundamental requirements for delayed logging in XFS are simple:
+
+	1. Reduce the amount of metadata written to the log by at least
+	   an order of magnitude.
+	2. Supply sufficient statistics to validate Requirement #1.
+	3. Supply sufficient new tracing infrastructure to be able to debug
+	   problems with the new code.
+	4. No on-disk format change (metadata or log format).
+	5. Enable and disable with a mount option.
+	6. No performance regressions for synchronous transaction workloads.
+
+Delayed Logging: Design
+-----------------------
+
+Storing Changes
+
+The problem with accumulating changes at a logical level (i.e. just using the
+existing log item dirty region tracking) is that when it comes to writing the
+changes to the log buffers, we need to ensure that the object we are formatting
+is not changing while we do this. This requires locking the object to prevent
+concurrent modification. Hence flushing the logical changes to the log would
+require us to lock every object, format them, and then unlock them again.
+
+This introduces lots of scope for deadlocks with transactions that are already
+running. For example, a transaction has object A locked and modified, but needs
+the delayed logging tracking lock to commit the transaction. However, the
+flushing thread has the delayed logging tracking lock already held, and is
+trying to get the lock on object A to flush it to the log buffer. This appears
+to be an unsolvable deadlock condition, and it was solving this problem that
+was the barrier to implementing delayed logging for so long.
+
+The solution is relatively simple - it just took a long time to recognise it.
+Put simply, the current logging code formats the changes to each item into an
+vector array that points to the changed regions in the item. The log write code
+simply copies the memory these vectors point to into the log buffer during
+transaction commit while the item is locked in the transaction. Instead of
+using the log buffer as the destination of the formatting code, we can use an
+allocated memory buffer big enough to fit the formatted vector.
+
+If we then copy the vector into the memory buffer and rewrite the vector to
+point to the memory buffer rather than the object itself, we now have a copy of
+the changes in a format that is compatible with the log buffer writing code.
+that does not require us to lock the item to access. This formatting and
+rewriting can all be done while the object is locked during transaction commit,
+resulting in a vector that is transactionally consistent and can be accessed
+without needing to lock the owning item.
+
+Hence we avoid the need to lock items when we need to flush outstanding
+asynchronous transactions to the log. The differences between the existing
+formatting method and the delayed logging formatting can be seen in the
+diagram below.
+
+Current format log vector:
+
+Object    +---------------------------------------------+
+Vector 1      +----+
+Vector 2                    +----+
+Vector 3                                   +----------+
+
+After formatting:
+
+Log Buffer    +-V1-+-V2-+----V3----+
+
+Delayed logging vector:
+
+Object    +---------------------------------------------+
+Vector 1      +----+
+Vector 2                    +----+
+Vector 3                                   +----------+
+
+After formatting:
+
+Memory Buffer +-V1-+-V2-+----V3----+
+Vector 1      +----+
+Vector 2           +----+
+Vector 3                +----------+
+
+The memory buffer and associated vector need to be passed as a single object,
+but still need to be associated with the parent object so if the object is
+relogged we can replace the current memory buffer with a new memory buffer that
+contains the latest changes.
+
+The reason for keeping the vector around after we've formatted the memory
+buffer is to support splitting vectors across log buffer boundaries correctly.
+If we don't keep the vector around, we do not know where the region boundaries
+are in the item, so we'd need a new encapsulation method for regions in the log
+buffer writing (i.e. double encapsulation). This would be an on-disk format
+change and as such is not desirable.  It also means we'd have to write the log
+region headers in the formatting stage, which is problematic as there is per
+region state that needs to be placed into the headers during the log write.
+
+Hence we need to keep the vector, but by attaching the memory buffer to it and
+rewriting the vector addresses to point at the memory buffer we end up with a
+self-describing object that can be passed to the log buffer write code to be
+handled in exactly the same manner as the existing log vectors are handled.
+Hence we avoid needing a new on-disk format to handle items that have been
+relogged in memory.
+
+
+Tracking Changes
+
+Now that we can record transactional changes in memory in a form that allows
+them to be used without limitations, we need to be able to track and accumulate
+them so that they can be written to the log at some later point in time.  The
+log item is the natural place to store this vector and buffer, and also makes sense
+to be the object that is used to track committed objects as it will always
+exist once the object has been included in a transaction.
+
+The log item is already used to track the log items that have been written to
+the log but not yet written to disk. Such log items are considered "active"
+and as such are stored in the Active Item List (AIL) which is a LSN-ordered
+double linked list. Items are inserted into this list during log buffer IO
+completion, after which they are unpinned and can be written to disk. An object
+that is in the AIL can be relogged, which causes the object to be pinned again
+and then moved forward in the AIL when the log buffer IO completes for that
+transaction.
+
+Essentially, this shows that an item that is in the AIL can still be modified
+and relogged, so any tracking must be separate to the AIL infrastructure. As
+such, we cannot reuse the AIL list pointers for tracking committed items, nor
+can we store state in any field that is protected by the AIL lock. Hence the
+committed item tracking needs it's own locks, lists and state fields in the log
+item.
+
+Similar to the AIL, tracking of committed items is done through a new list
+called the Committed Item List (CIL).  The list tracks log items that have been
+committed and have formatted memory buffers attached to them. It tracks objects
+in transaction commit order, so when an object is relogged it is removed from
+it's place in the list and re-inserted at the tail. This is entirely arbitrary
+and done to make it easy for debugging - the last items in the list are the
+ones that are most recently modified. Ordering of the CIL is not necessary for
+transactional integrity (as discussed in the next section) so the ordering is
+done for convenience/sanity of the developers.
+
+
+Delayed Logging: Checkpoints
+
+When we have a log synchronisation event, commonly known as a "log force",
+all the items in the CIL must be written into the log via the log buffers.
+We need to write these items in the order that they exist in the CIL, and they
+need to be written as an atomic transaction. The need for all the objects to be
+written as an atomic transaction comes from the requirements of relogging and
+log replay - all the changes in all the objects in a given transaction must
+either be completely replayed during log recovery, or not replayed at all. If
+a transaction is not replayed because it is not complete in the log, then
+no later transactions should be replayed, either.
+
+To fulfill this requirement, we need to write the entire CIL in a single log
+transaction. Fortunately, the XFS log code has no fixed limit on the size of a
+transaction, nor does the log replay code. The only fundamental limit is that
+the transaction cannot be larger than just under half the size of the log.  The
+reason for this limit is that to find the head and tail of the log, there must
+be at least one complete transaction in the log at any given time. If a
+transaction is larger than half the log, then there is the possibility that a
+crash during the write of a such a transaction could partially overwrite the
+only complete previous transaction in the log. This will result in a recovery
+failure and an inconsistent filesystem and hence we must enforce the maximum
+size of a checkpoint to be slightly less than a half the log.
+
+Apart from this size requirement, a checkpoint transaction looks no different
+to any other transaction - it contains a transaction header, a series of
+formatted log items and a commit record at the tail. From a recovery
+perspective, the checkpoint transaction is also no different - just a lot
+bigger with a lot more items in it. The worst case effect of this is that we
+might need to tune the recovery transaction object hash size.
+
+Because the checkpoint is just another transaction and all the changes to log
+items are stored as log vectors, we can use the existing log buffer writing
+code to write the changes into the log. To do this efficiently, we need to
+minimise the time we hold the CIL locked while writing the checkpoint
+transaction. The current log write code enables us to do this easily with the
+way it separates the writing of the transaction contents (the log vectors) from
+the transaction commit record, but tracking this requires us to have a
+per-checkpoint context that travels through the log write process through to
+checkpoint completion.
+
+Hence a checkpoint has a context that tracks the state of the current
+checkpoint from initiation to checkpoint completion. A new context is initiated
+at the same time a checkpoint transaction is started. That is, when we remove
+all the current items from the CIL during a checkpoint operation, we move all
+those changes into the current checkpoint context. We then initialise a new
+context and attach that to the CIL for aggregation of new transactions.
+
+This allows us to unlock the CIL immediately after transfer of all the
+committed items and effectively allow new transactions to be issued while we
+are formatting the checkpoint into the log. It also allows concurrent
+checkpoints to be written into the log buffers in the case of log force heavy
+workloads, just like the existing transaction commit code does. This, however,
+requires that we strictly order the commit records in the log so that
+checkpoint sequence order is maintained during log replay.
+
+To ensure that we can be writing an item into a checkpoint transaction at
+the same time another transaction modifies the item and inserts the log item
+into the new CIL, then checkpoint transaction commit code cannot use log items
+to store the list of log vectors that need to be written into the transaction.
+Hence log vectors need to be able to be chained together to allow them to be
+detached from the log items. That is, when the CIL is flushed the memory
+buffer and log vector attached to each log item needs to be attached to the
+checkpoint context so that the log item can be released. In diagrammatic form,
+the CIL would look like this before the flush:
+
+	CIL Head
+	   |
+	   V
+	Log Item <-> log vector 1	-> memory buffer
+	   |				-> vector array
+	   V
+	Log Item <-> log vector 2	-> memory buffer
+	   |				-> vector array
+	   V
+	......
+	   |
+	   V
+	Log Item <-> log vector N-1	-> memory buffer
+	   |				-> vector array
+	   V
+	Log Item <-> log vector N	-> memory buffer
+					-> vector array
+
+And after the flush the CIL head is empty, and the checkpoint context log
+vector list would look like:
+
+	Checkpoint Context
+	   |
+	   V
+	log vector 1	-> memory buffer
+	   |		-> vector array
+	   |		-> Log Item
+	   V
+	log vector 2	-> memory buffer
+	   |		-> vector array
+	   |		-> Log Item
+	   V
+	......
+	   |
+	   V
+	log vector N-1	-> memory buffer
+	   |		-> vector array
+	   |		-> Log Item
+	   V
+	log vector N	-> memory buffer
+			-> vector array
+			-> Log Item
+
+Once this transfer is done, the CIL can be unlocked and new transactions can
+start, while the checkpoint flush code works over the log vector chain to
+commit the checkpoint.
+
+Once the checkpoint is written into the log buffers, the checkpoint context is
+attached to the log buffer that the commit record was written to along with a
+completion callback. Log IO completion will call that callback, which can then
+run transaction committed processing for the log items (i.e. insert into AIL
+and unpin) in the log vector chain and then free the log vector chain and
+checkpoint context.
+
+Discussion Point: I am uncertain as to whether the log item is the most
+efficient way to track vectors, even though it seems like the natural way to do
+it. The fact that we walk the log items (in the CIL) just to chain the log
+vectors and break the link between the log item and the log vector means that
+we take a cache line hit for the log item list modification, then another for
+the log vector chaining. If we track by the log vectors, then we only need to
+break the link between the log item and the log vector, which means we should
+dirty only the log item cachelines. Normally I wouldn't be concerned about one
+vs two dirty cachelines except for the fact I've seen upwards of 80,000 log
+vectors in one checkpoint transaction. I'd guess this is a "measure and
+compare" situation that can be done after a working and reviewed implementation
+is in the dev tree....
+
+Delayed Logging: Checkpoint Sequencing
+
+One of the key aspects of the XFS transaction subsystem is that it tags
+committed transactions with the log sequence number of the transaction commit.
+This allows transactions to be issued asynchronously even though there may be
+future operations that cannot be completed until that transaction is fully
+committed to the log. In the rare case that a dependent operation occurs (e.g.
+re-using a freed metadata extent for a data extent), a special, optimised log
+force can be issued to force the dependent transaction to disk immediately.
+
+To do this, transactions need to record the LSN of the commit record of the
+transaction. This LSN comes directly from the log buffer the transaction is
+written into. While this works just fine for the existing transaction
+mechanism, it does not work for delayed logging because transactions are not
+written directly into the log buffers. Hence some other method of sequencing
+transactions is required.
+
+As discussed in the checkpoint section, delayed logging uses per-checkpoint
+contexts, and as such it is simple to assign a sequence number to each
+checkpoint. Because the switching of checkpoint contexts must be done
+atomically, it is simple to ensure that each new context has a monotonically
+increasing sequence number assigned to it without the need for an external
+atomic counter - we can just take the current context sequence number and add
+one to it for the new context.
+
+Then, instead of assigning a log buffer LSN to the transaction commit LSN
+during the commit, we can assign the current checkpoint sequence. This allows
+operations that track transactions that have not yet completed know what
+checkpoint sequence needs to be committed before they can continue. As a
+result, the code that forces the log to a specific LSN now needs to ensure that
+the log forces to a specific checkpoint.
+
+To ensure that we can do this, we need to track all the checkpoint contexts
+that are currently committing to the log. When we flush a checkpoint, the
+context gets added to a "committing" list which can be searched. When a
+checkpoint commit completes, it is removed from the committing list. Because
+the checkpoint context records the LSN of the commit record for the checkpoint,
+we can also wait on the log buffer that contains the commit record, thereby
+using the existing log force mechanisms to execute synchronous forces.
+
+It should be noted that the synchronous forces may need to be extended with
+mitigation algorithms similar to the current log buffer code to allow
+aggregation of multiple synchronous transactions if there are already
+synchronous transactions being flushed. Investigation of the performance of the
+current design is needed before making any decisions here.
+
+The main concern with log forces is to ensure that all the previous checkpoints
+are also committed to disk before the one we need to wait for. Therefore we
+need to check that all the prior contexts in the committing list are also
+complete before waiting on the one we need to complete. We do this
+synchronisation in the log force code so that we don't need to wait anywhere
+else for such serialisation - it only matters when we do a log force.
+
+The only remaining complexity is that a log force now also has to handle the
+case where the forcing sequence number is the same as the current context. That
+is, we need to flush the CIL and potentially wait for it to complete. This is a
+simple addition to the existing log forcing code to check the sequence numbers
+and push if required. Indeed, placing the current sequence checkpoint flush in
+the log force code enables the current mechanism for issuing synchronous
+transactions to remain untouched (i.e. commit an asynchronous transaction, then
+force the log at the LSN of that transaction) and so the higher level code
+behaves the same regardless of whether delayed logging is being used or not.
+
+Delayed Logging: Checkpoint Log Space Accounting
+
+The big issue for a checkpoint transaction is the log space reservation for the
+transaction. We don't know how big a checkpoint transaction is going to be
+ahead of time, nor how many log buffers it will take to write out, nor the
+number of split log vector regions are going to be used. We can track the
+amount of log space required as we add items to the commit item list, but we
+still need to reserve the space in the log for the checkpoint.
+
+A typical transaction reserves enough space in the log for the worst case space
+usage of the transaction. The reservation accounts for log record headers,
+transaction and region headers, headers for split regions, buffer tail padding,
+etc. as well as the actual space for all the changed metadata in the
+transaction. While some of this is fixed overhead, much of it is dependent on
+the size of the transaction and the number of regions being logged (the number
+of log vectors in the transaction).
+
+An example of the differences would be logging directory changes versus logging
+inode changes. If you modify lots of inode cores (e.g. chmod -R g+w *), then
+there are lots of transactions that only contain an inode core and an inode log
+format structure. That is, two vectors totaling roughly 150 bytes. If we modify
+10,000 inodes, we have about 1.5MB of metadata to write in 20,000 vectors. Each
+vector is 12 bytes, so the total to be logged is approximately 1.75MB. In
+comparison, if we are logging full directory buffers, they are typically 4KB
+each, so we in 1.5MB of directory buffers we'd have roughly 400 buffers and a
+buffer format structure for each buffer - roughly 800 vectors or 1.51MB total
+space.  From this, it should be obvious that a static log space reservation is
+not particularly flexible and is difficult to select the "optimal value" for
+all workloads.
+
+Further, if we are going to use a static reservation, which bit of the entire
+reservation does it cover? We account for space used by the transaction
+reservation by tracking the space currently used by the object in the CIL and
+then calculating the increase or decrease in space used as the object is
+relogged. This allows for a checkpoint reservation to only have to account for
+log buffer metadata used such as log header records.
+
+However, even using a static reservation for just the log metadata is
+problematic. Typically log record headers use at least 16KB of log space per
+1MB of log space consumed (512 bytes per 32k) and the reservation needs to be
+large enough to handle arbitrary sized checkpoint transactions. This
+reservation needs to be made before the checkpoint is started, and we need to
+be able to reserve the space without sleeping.  For a 8MB checkpoint, we need a
+reservation of around 150KB, which is a non-trivial amount of space.
+
+A static reservation needs to manipulate the log grant counters - we can take a
+permanent reservation on the space, but we still need to make sure we refresh
+the write reservation (the actual space available to the transaction) after
+every checkpoint transaction completion. Unfortunately, if this space is not
+available when required, then the regrant code will sleep waiting for it.
+
+The problem with this is that it can lead to deadlocks as we may need to commit
+checkpoints to be able to free up log space (refer back to the description of
+rolling transactions for an example of this).  Hence we *must* always have
+space available in the log if we are to use static reservations, and that is
+very difficult and complex to arrange. It is possible to do, but there is a
+simpler way.
+
+The simpler way of doing this is tracking the entire log space used by the
+items in the CIL and using this to dynamically calculate the amount of log
+space required by the log metadata. If this log metadata space changes as a
+result of a transaction commit inserting a new memory buffer into the CIL, then
+the difference in space required is removed from the transaction that causes
+the change. Transactions at this level will *always* have enough space
+available in their reservation for this as they have already reserved the
+maximal amount of log metadata space they require, and such a delta reservation
+will always be less than or equal to the maximal amount in the reservation.
+
+Hence we can grow the checkpoint transaction reservation dynamically as items
+are added to the CIL and avoid the need for reserving and regranting log space
+up front. This avoids deadlocks and removes a blocking point from the
+checkpoint flush code.
+
+As mentioned early, transactions can't grow to more than half the size of the
+log. Hence as part of the reservation growing, we need to also check the size
+of the reservation against the maximum allowed transaction size. If we reach
+the maximum threshold, we need to push the CIL to the log. This is effectively
+a "background flush" and is done on demand. This is identical to
+a CIL push triggered by a log force, only that there is no waiting for the
+checkpoint commit to complete. This background push is checked and executed by
+transaction commit code.
+
+If the transaction subsystem goes idle while we still have items in the CIL,
+they will be flushed by the periodic log force issued by the xfssyncd. This log
+force will push the CIL to disk, and if the transaction subsystem stays idle,
+allow the idle log to be covered (effectively marked clean) in exactly the same
+manner that is done for the existing logging method. A discussion point is
+whether this log force needs to be done more frequently than the current rate
+which is once every 30s.
+
+
+Delayed Logging: Log Item Pinning
+
+Currently log items are pinned during transaction commit while the items are
+still locked. This happens just after the items are formatted, though it could
+be done any time before the items are unlocked. The result of this mechanism is
+that items get pinned once for every transaction that is committed to the log
+buffers. Hence items that are relogged in the log buffers will have a pin count
+for every outstanding transaction they were dirtied in. When each of these
+transactions is completed, they will unpin the item once. As a result, the item
+only becomes unpinned when all the transactions complete and there are no
+pending transactions. Thus the pinning and unpinning of a log item is symmetric
+as there is a 1:1 relationship with transaction commit and log item completion.
+
+For delayed logging, however, we have an asymmetric transaction commit to
+completion relationship. Every time an object is relogged in the CIL it goes
+through the commit process without a corresponding completion being registered.
+That is, we now have a many-to-one relationship between transaction commit and
+log item completion. The result of this is that pinning and unpinning of the
+log items becomes unbalanced if we retain the "pin on transaction commit, unpin
+on transaction completion" model.
+
+To keep pin/unpin symmetry, the algorithm needs to change to a "pin on
+insertion into the CIL, unpin on checkpoint completion". In other words, the
+pinning and unpinning becomes symmetric around a checkpoint context. We have to
+pin the object the first time it is inserted into the CIL - if it is already in
+the CIL during a transaction commit, then we do not pin it again. Because there
+can be multiple outstanding checkpoint contexts, we can still see elevated pin
+counts, but as each checkpoint completes the pin count will retain the correct
+value according to it's context.
+
+Just to make matters more slightly more complex, this checkpoint level context
+for the pin count means that the pinning of an item must take place under the
+CIL commit/flush lock. If we pin the object outside this lock, we cannot
+guarantee which context the pin count is associated with. This is because of
+the fact pinning the item is dependent on whether the item is present in the
+current CIL or not. If we don't pin the CIL first before we check and pin the
+object, we have a race with CIL being flushed between the check and the pin
+(or not pinning, as the case may be). Hence we must hold the CIL flush/commit
+lock to guarantee that we pin the items correctly.
+
+Delayed Logging: Concurrent Scalability
+
+A fundamental requirement for the CIL is that accesses through transaction
+commits must scale to many concurrent commits. The current transaction commit
+code does not break down even when there are transactions coming from 2048
+processors at once. The current transaction code does not go any faster than if
+there was only one CPU using it, but it does not slow down either.
+
+As a result, the delayed logging transaction commit code needs to be designed
+for concurrency from the ground up. It is obvious that there are serialisation
+points in the design - the three important ones are:
+
+	1. Locking out new transaction commits while flushing the CIL
+	2. Adding items to the CIL and updating item space accounting
+	3. Checkpoint commit ordering
+
+Looking at the transaction commit and CIL flushing interactions, it is clear
+that we have a many-to-one interaction here. That is, the only restriction on
+the number of concurrent transactions that can be trying to commit at once is
+the amount of space available in the log for their reservations. The practical
+limit here is in the order of several hundred concurrent transactions for a
+128MB log, which means that it is generally one per CPU in a machine.
+
+The amount of time a transaction commit needs to hold out a flush is a
+relatively long period of time - the pinning of log items needs to be done
+while we are holding out a CIL flush, so at the moment that means it is held
+across the formatting of the objects into memory buffers (i.e. while memcpy()s
+are in progress). Ultimately a two pass algorithm where the formatting is done
+separately to the pinning of objects could be used to reduce the hold time of
+the transaction commit side.
+
+Because of the number of potential transaction commit side holders, the lock
+really needs to be a sleeping lock - if the CIL flush takes the lock, we do not
+want every other CPU in the machine spinning on the CIL lock. Given that
+flushing the CIL could involve walking a list of tens of thousands of log
+items, it will get held for a significant time and so spin contention is a
+significant concern. Preventing lots of CPUs spinning doing nothing is the
+main reason for choosing a sleeping lock even though nothing in either the
+transaction commit or CIL flush side sleeps with the lock held.
+
+It should also be noted that CIL flushing is also a relatively rare operation
+compared to transaction commit for asynchronous transaction workloads - only
+time will tell if using a read-write semaphore for exclusion will limit
+transaction commit concurrency due to cache line bouncing of the lock on the
+read side.
+
+The second serialisation point is on the transaction commit side where items
+are inserted into the CIL. Because transactions can enter this code
+concurrently, the CIL needs to be protected separately from the above
+commit/flush exclusion. It also needs to be an exclusive lock but it is only
+held for a very short time and so a spin lock is appropriate here. It is
+possible that this lock will become a contention point, but given the short
+hold time once per transaction I think that contention is unlikely.
+
+The final serialisation point is the checkpoint commit record ordering code
+that is run as part of the checkpoint commit and log force sequencing. The code
+path that triggers a CIL flush (i.e. whatever triggers the log force) will enter
+an ordering loop after writing all the log vectors into the log buffers but
+before writing the commit record. This loop walks the list of committing
+checkpoints and needs to block waiting for checkpoints to complete their commit
+record write. As a result it needs a lock and a wait variable. Log force
+sequencing also requires the same lock, list walk, and blocking mechanism to
+ensure completion of checkpoints.
+
+These two sequencing operations can use the mechanism even though the
+events they are waiting for are different. The checkpoint commit record
+sequencing needs to wait until checkpoint contexts contain a commit LSN
+(obtained through completion of a commit record write) while log force
+sequencing needs to wait until previous checkpoint contexts are removed from
+the committing list (i.e. they've completed). A simple wait variable and
+broadcast wakeups (thundering herds) has been used to implement these two
+serialisation queues. They use the same lock as the CIL, too. If we see too
+much contention on the CIL lock, or too many context switches as a result of
+the broadcast wakeups these operations can be put under a new spinlock and
+given separate wait lists to reduce lock contention and the number of processes
+woken by the wrong event.
+
+
+Lifecycle Changes
+
+The existing log item life cycle is as follows:
+
+	1. Transaction allocate
+	2. Transaction reserve
+	3. Lock item
+	4. Join item to transaction
+		If not already attached,
+			Allocate log item
+			Attach log item to owner item
+		Attach log item to transaction
+	5. Modify item
+		Record modifications in log item
+	6. Transaction commit
+		Pin item in memory
+		Format item into log buffer
+		Write commit LSN into transaction
+		Unlock item
+		Attach transaction to log buffer
+
+	<log buffer IO dispatched>
+	<log buffer IO completes>
+
+	7. Transaction completion
+		Mark log item committed
+		Insert log item into AIL
+			Write commit LSN into log item
+		Unpin log item
+	8. AIL traversal
+		Lock item
+		Mark log item clean
+		Flush item to disk
+
+	<item IO completion>
+
+	9. Log item removed from AIL
+		Moves log tail
+		Item unlocked
+
+Essentially, steps 1-6 operate independently from step 7, which is also
+independent of steps 8-9. An item can be locked in steps 1-6 or steps 8-9
+at the same time step 7 is occurring, but only steps 1-6 or 8-9 can occur
+at the same time. If the log item is in the AIL or between steps 6 and 7
+and steps 1-6 are re-entered, then the item is relogged. Only when steps 8-9
+are entered and completed is the object considered clean.
+
+With delayed logging, there are new steps inserted into the life cycle:
+
+	1. Transaction allocate
+	2. Transaction reserve
+	3. Lock item
+	4. Join item to transaction
+		If not already attached,
+			Allocate log item
+			Attach log item to owner item
+		Attach log item to transaction
+	5. Modify item
+		Record modifications in log item
+	6. Transaction commit
+		Pin item in memory if not pinned in CIL
+		Format item into log vector + buffer
+		Attach log vector and buffer to log item
+		Insert log item into CIL
+		Write CIL context sequence into transaction
+		Unlock item
+
+	<next log force>
+
+	7. CIL push
+		lock CIL flush
+		Chain log vectors and buffers together
+		Remove items from CIL
+		unlock CIL flush
+		write log vectors into log
+		sequence commit records
+		attach checkpoint context to log buffer
+
+	<log buffer IO dispatched>
+	<log buffer IO completes>
+
+	8. Checkpoint completion
+		Mark log item committed
+		Insert item into AIL
+			Write commit LSN into log item
+		Unpin log item
+	9. AIL traversal
+		Lock item
+		Mark log item clean
+		Flush item to disk
+	<item IO completion>
+	10. Log item removed from AIL
+		Moves log tail
+		Item unlocked
+
+From this, it can be seen that the only life cycle differences between the two
+logging methods are in the middle of the life cycle - they still have the same
+beginning and end and execution constraints. The only differences are in the
+committing of the log items to the log itself and the completion processing.
+Hence delayed logging should not introduce any constraints on log item
+behaviour, allocation or freeing that don't already exist.
+
+As a result of this zero-impact "insertion" of delayed logging infrastructure
+and the design of the internal structures to avoid on disk format changes, we
+can basically switch between delayed logging and the existing mechanism with a
+mount option. Fundamentally, there is no reason why the log manager would not
+be able to swap methods automatically and transparently depending on load
+characteristics, but this should not be necessary if delayed logging works as
+designed.
diff --git a/Documentation/filesystems/xfs-self-describing-metadata.txt b/Documentation/filesystems/xfs-self-describing-metadata.txt
new file mode 100644
index 000000000..05aa45516
--- /dev/null
+++ b/Documentation/filesystems/xfs-self-describing-metadata.txt
@@ -0,0 +1,350 @@
+XFS Self Describing Metadata
+----------------------------
+
+Introduction
+------------
+
+The largest scalability problem facing XFS is not one of algorithmic
+scalability, but of verification of the filesystem structure. Scalabilty of the
+structures and indexes on disk and the algorithms for iterating them are
+adequate for supporting PB scale filesystems with billions of inodes, however it
+is this very scalability that causes the verification problem.
+
+Almost all metadata on XFS is dynamically allocated. The only fixed location
+metadata is the allocation group headers (SB, AGF, AGFL and AGI), while all
+other metadata structures need to be discovered by walking the filesystem
+structure in different ways. While this is already done by userspace tools for
+validating and repairing the structure, there are limits to what they can
+verify, and this in turn limits the supportable size of an XFS filesystem.
+
+For example, it is entirely possible to manually use xfs_db and a bit of
+scripting to analyse the structure of a 100TB filesystem when trying to
+determine the root cause of a corruption problem, but it is still mainly a
+manual task of verifying that things like single bit errors or misplaced writes
+weren't the ultimate cause of a corruption event. It may take a few hours to a
+few days to perform such forensic analysis, so for at this scale root cause
+analysis is entirely possible.
+
+However, if we scale the filesystem up to 1PB, we now have 10x as much metadata
+to analyse and so that analysis blows out towards weeks/months of forensic work.
+Most of the analysis work is slow and tedious, so as the amount of analysis goes
+up, the more likely that the cause will be lost in the noise.  Hence the primary
+concern for supporting PB scale filesystems is minimising the time and effort
+required for basic forensic analysis of the filesystem structure.
+
+
+Self Describing Metadata
+------------------------
+
+One of the problems with the current metadata format is that apart from the
+magic number in the metadata block, we have no other way of identifying what it
+is supposed to be. We can't even identify if it is the right place. Put simply,
+you can't look at a single metadata block in isolation and say "yes, it is
+supposed to be there and the contents are valid".
+
+Hence most of the time spent on forensic analysis is spent doing basic
+verification of metadata values, looking for values that are in range (and hence
+not detected by automated verification checks) but are not correct. Finding and
+understanding how things like cross linked block lists (e.g. sibling
+pointers in a btree end up with loops in them) are the key to understanding what
+went wrong, but it is impossible to tell what order the blocks were linked into
+each other or written to disk after the fact.
+
+Hence we need to record more information into the metadata to allow us to
+quickly determine if the metadata is intact and can be ignored for the purpose
+of analysis. We can't protect against every possible type of error, but we can
+ensure that common types of errors are easily detectable.  Hence the concept of
+self describing metadata.
+
+The first, fundamental requirement of self describing metadata is that the
+metadata object contains some form of unique identifier in a well known
+location. This allows us to identify the expected contents of the block and
+hence parse and verify the metadata object. IF we can't independently identify
+the type of metadata in the object, then the metadata doesn't describe itself
+very well at all!
+
+Luckily, almost all XFS metadata has magic numbers embedded already - only the
+AGFL, remote symlinks and remote attribute blocks do not contain identifying
+magic numbers. Hence we can change the on-disk format of all these objects to
+add more identifying information and detect this simply by changing the magic
+numbers in the metadata objects. That is, if it has the current magic number,
+the metadata isn't self identifying. If it contains a new magic number, it is
+self identifying and we can do much more expansive automated verification of the
+metadata object at runtime, during forensic analysis or repair.
+
+As a primary concern, self describing metadata needs some form of overall
+integrity checking. We cannot trust the metadata if we cannot verify that it has
+not been changed as a result of external influences. Hence we need some form of
+integrity check, and this is done by adding CRC32c validation to the metadata
+block. If we can verify the block contains the metadata it was intended to
+contain, a large amount of the manual verification work can be skipped.
+
+CRC32c was selected as metadata cannot be more than 64k in length in XFS and
+hence a 32 bit CRC is more than sufficient to detect multi-bit errors in
+metadata blocks. CRC32c is also now hardware accelerated on common CPUs so it is
+fast. So while CRC32c is not the strongest of possible integrity checks that
+could be used, it is more than sufficient for our needs and has relatively
+little overhead. Adding support for larger integrity fields and/or algorithms
+does really provide any extra value over CRC32c, but it does add a lot of
+complexity and so there is no provision for changing the integrity checking
+mechanism.
+
+Self describing metadata needs to contain enough information so that the
+metadata block can be verified as being in the correct place without needing to
+look at any other metadata. This means it needs to contain location information.
+Just adding a block number to the metadata is not sufficient to protect against
+mis-directed writes - a write might be misdirected to the wrong LUN and so be
+written to the "correct block" of the wrong filesystem. Hence location
+information must contain a filesystem identifier as well as a block number.
+
+Another key information point in forensic analysis is knowing who the metadata
+block belongs to. We already know the type, the location, that it is valid
+and/or corrupted, and how long ago that it was last modified. Knowing the owner
+of the block is important as it allows us to find other related metadata to
+determine the scope of the corruption. For example, if we have a extent btree
+object, we don't know what inode it belongs to and hence have to walk the entire
+filesystem to find the owner of the block. Worse, the corruption could mean that
+no owner can be found (i.e. it's an orphan block), and so without an owner field
+in the metadata we have no idea of the scope of the corruption. If we have an
+owner field in the metadata object, we can immediately do top down validation to
+determine the scope of the problem.
+
+Different types of metadata have different owner identifiers. For example,
+directory, attribute and extent tree blocks are all owned by an inode, whilst
+freespace btree blocks are owned by an allocation group. Hence the size and
+contents of the owner field are determined by the type of metadata object we are
+looking at.  The owner information can also identify misplaced writes (e.g.
+freespace btree block written to the wrong AG).
+
+Self describing metadata also needs to contain some indication of when it was
+written to the filesystem. One of the key information points when doing forensic
+analysis is how recently the block was modified. Correlation of set of corrupted
+metadata blocks based on modification times is important as it can indicate
+whether the corruptions are related, whether there's been multiple corruption
+events that lead to the eventual failure, and even whether there are corruptions
+present that the run-time verification is not detecting.
+
+For example, we can determine whether a metadata object is supposed to be free
+space or still allocated if it is still referenced by its owner by looking at
+when the free space btree block that contains the block was last written
+compared to when the metadata object itself was last written.  If the free space
+block is more recent than the object and the object's owner, then there is a
+very good chance that the block should have been removed from the owner.
+
+To provide this "written timestamp", each metadata block gets the Log Sequence
+Number (LSN) of the most recent transaction it was modified on written into it.
+This number will always increase over the life of the filesystem, and the only
+thing that resets it is running xfs_repair on the filesystem. Further, by use of
+the LSN we can tell if the corrupted metadata all belonged to the same log
+checkpoint and hence have some idea of how much modification occurred between
+the first and last instance of corrupt metadata on disk and, further, how much
+modification occurred between the corruption being written and when it was
+detected.
+
+Runtime Validation
+------------------
+
+Validation of self-describing metadata takes place at runtime in two places:
+
+	- immediately after a successful read from disk
+	- immediately prior to write IO submission
+
+The verification is completely stateless - it is done independently of the
+modification process, and seeks only to check that the metadata is what it says
+it is and that the metadata fields are within bounds and internally consistent.
+As such, we cannot catch all types of corruption that can occur within a block
+as there may be certain limitations that operational state enforces of the
+metadata, or there may be corruption of interblock relationships (e.g. corrupted
+sibling pointer lists). Hence we still need stateful checking in the main code
+body, but in general most of the per-field validation is handled by the
+verifiers.
+
+For read verification, the caller needs to specify the expected type of metadata
+that it should see, and the IO completion process verifies that the metadata
+object matches what was expected. If the verification process fails, then it
+marks the object being read as EFSCORRUPTED. The caller needs to catch this
+error (same as for IO errors), and if it needs to take special action due to a
+verification error it can do so by catching the EFSCORRUPTED error value. If we
+need more discrimination of error type at higher levels, we can define new
+error numbers for different errors as necessary.
+
+The first step in read verification is checking the magic number and determining
+whether CRC validating is necessary. If it is, the CRC32c is calculated and
+compared against the value stored in the object itself. Once this is validated,
+further checks are made against the location information, followed by extensive
+object specific metadata validation. If any of these checks fail, then the
+buffer is considered corrupt and the EFSCORRUPTED error is set appropriately.
+
+Write verification is the opposite of the read verification - first the object
+is extensively verified and if it is OK we then update the LSN from the last
+modification made to the object, After this, we calculate the CRC and insert it
+into the object. Once this is done the write IO is allowed to continue. If any
+error occurs during this process, the buffer is again marked with a EFSCORRUPTED
+error for the higher layers to catch.
+
+Structures
+----------
+
+A typical on-disk structure needs to contain the following information:
+
+struct xfs_ondisk_hdr {
+        __be32  magic;		/* magic number */
+        __be32  crc;		/* CRC, not logged */
+        uuid_t  uuid;		/* filesystem identifier */
+        __be64  owner;		/* parent object */
+        __be64  blkno;		/* location on disk */
+        __be64  lsn;		/* last modification in log, not logged */
+};
+
+Depending on the metadata, this information may be part of a header structure
+separate to the metadata contents, or may be distributed through an existing
+structure. The latter occurs with metadata that already contains some of this
+information, such as the superblock and AG headers.
+
+Other metadata may have different formats for the information, but the same
+level of information is generally provided. For example:
+
+	- short btree blocks have a 32 bit owner (ag number) and a 32 bit block
+	  number for location. The two of these combined provide the same
+	  information as @owner and @blkno in eh above structure, but using 8
+	  bytes less space on disk.
+
+	- directory/attribute node blocks have a 16 bit magic number, and the
+	  header that contains the magic number has other information in it as
+	  well. hence the additional metadata headers change the overall format
+	  of the metadata.
+
+A typical buffer read verifier is structured as follows:
+
+#define XFS_FOO_CRC_OFF		offsetof(struct xfs_ondisk_hdr, crc)
+
+static void
+xfs_foo_read_verify(
+	struct xfs_buf	*bp)
+{
+       struct xfs_mount *mp = bp->b_target->bt_mount;
+
+        if ((xfs_sb_version_hascrc(&mp->m_sb) &&
+             !xfs_verify_cksum(bp->b_addr, BBTOB(bp->b_length),
+					XFS_FOO_CRC_OFF)) ||
+            !xfs_foo_verify(bp)) {
+                XFS_CORRUPTION_ERROR(__func__, XFS_ERRLEVEL_LOW, mp, bp->b_addr);
+                xfs_buf_ioerror(bp, EFSCORRUPTED);
+        }
+}
+
+The code ensures that the CRC is only checked if the filesystem has CRCs enabled
+by checking the superblock of the feature bit, and then if the CRC verifies OK
+(or is not needed) it verifies the actual contents of the block.
+
+The verifier function will take a couple of different forms, depending on
+whether the magic number can be used to determine the format of the block. In
+the case it can't, the code is structured as follows:
+
+static bool
+xfs_foo_verify(
+	struct xfs_buf		*bp)
+{
+        struct xfs_mount	*mp = bp->b_target->bt_mount;
+        struct xfs_ondisk_hdr	*hdr = bp->b_addr;
+
+        if (hdr->magic != cpu_to_be32(XFS_FOO_MAGIC))
+                return false;
+
+        if (!xfs_sb_version_hascrc(&mp->m_sb)) {
+		if (!uuid_equal(&hdr->uuid, &mp->m_sb.sb_uuid))
+			return false;
+		if (bp->b_bn != be64_to_cpu(hdr->blkno))
+			return false;
+		if (hdr->owner == 0)
+			return false;
+	}
+
+	/* object specific verification checks here */
+
+        return true;
+}
+
+If there are different magic numbers for the different formats, the verifier
+will look like:
+
+static bool
+xfs_foo_verify(
+	struct xfs_buf		*bp)
+{
+        struct xfs_mount	*mp = bp->b_target->bt_mount;
+        struct xfs_ondisk_hdr	*hdr = bp->b_addr;
+
+        if (hdr->magic == cpu_to_be32(XFS_FOO_CRC_MAGIC)) {
+		if (!uuid_equal(&hdr->uuid, &mp->m_sb.sb_uuid))
+			return false;
+		if (bp->b_bn != be64_to_cpu(hdr->blkno))
+			return false;
+		if (hdr->owner == 0)
+			return false;
+	} else if (hdr->magic != cpu_to_be32(XFS_FOO_MAGIC))
+		return false;
+
+	/* object specific verification checks here */
+
+        return true;
+}
+
+Write verifiers are very similar to the read verifiers, they just do things in
+the opposite order to the read verifiers. A typical write verifier:
+
+static void
+xfs_foo_write_verify(
+	struct xfs_buf	*bp)
+{
+	struct xfs_mount	*mp = bp->b_target->bt_mount;
+	struct xfs_buf_log_item	*bip = bp->b_fspriv;
+
+	if (!xfs_foo_verify(bp)) {
+		XFS_CORRUPTION_ERROR(__func__, XFS_ERRLEVEL_LOW, mp, bp->b_addr);
+		xfs_buf_ioerror(bp, EFSCORRUPTED);
+		return;
+	}
+
+	if (!xfs_sb_version_hascrc(&mp->m_sb))
+		return;
+
+
+	if (bip) {
+		struct xfs_ondisk_hdr	*hdr = bp->b_addr;
+		hdr->lsn = cpu_to_be64(bip->bli_item.li_lsn);
+	}
+	xfs_update_cksum(bp->b_addr, BBTOB(bp->b_length), XFS_FOO_CRC_OFF);
+}
+
+This will verify the internal structure of the metadata before we go any
+further, detecting corruptions that have occurred as the metadata has been
+modified in memory. If the metadata verifies OK, and CRCs are enabled, we then
+update the LSN field (when it was last modified) and calculate the CRC on the
+metadata. Once this is done, we can issue the IO.
+
+Inodes and Dquots
+-----------------
+
+Inodes and dquots are special snowflakes. They have per-object CRC and
+self-identifiers, but they are packed so that there are multiple objects per
+buffer. Hence we do not use per-buffer verifiers to do the work of per-object
+verification and CRC calculations. The per-buffer verifiers simply perform basic
+identification of the buffer - that they contain inodes or dquots, and that
+there are magic numbers in all the expected spots. All further CRC and
+verification checks are done when each inode is read from or written back to the
+buffer.
+
+The structure of the verifiers and the identifiers checks is very similar to the
+buffer code described above. The only difference is where they are called. For
+example, inode read verification is done in xfs_iread() when the inode is first
+read out of the buffer and the struct xfs_inode is instantiated. The inode is
+already extensively verified during writeback in xfs_iflush_int, so the only
+addition here is to add the LSN and CRC to the inode as it is copied back into
+the buffer.
+
+XXX: inode unlinked list modification doesn't recalculate the inode CRC! None of
+the unlinked list modifications check or update CRCs, neither during unlink nor
+log recovery. So, it's gone unnoticed until now. This won't matter immediately -
+repair will probably complain about it - but it needs to be fixed.
+
diff --git a/Documentation/filesystems/xfs.txt b/Documentation/filesystems/xfs.txt
new file mode 100644
index 000000000..a9ae82fb9
--- /dev/null
+++ b/Documentation/filesystems/xfs.txt
@@ -0,0 +1,469 @@
+
+The SGI XFS Filesystem
+======================
+
+XFS is a high performance journaling filesystem which originated
+on the SGI IRIX platform.  It is completely multi-threaded, can
+support large files and large filesystems, extended attributes,
+variable block sizes, is extent based, and makes extensive use of
+Btrees (directories, extents, free space) to aid both performance
+and scalability.
+
+Refer to the documentation at https://xfs.wiki.kernel.org/
+for further details.  This implementation is on-disk compatible
+with the IRIX version of XFS.
+
+
+Mount Options
+=============
+
+When mounting an XFS filesystem, the following options are accepted.
+For boolean mount options, the names with the (*) suffix is the
+default behaviour.
+
+  allocsize=size
+	Sets the buffered I/O end-of-file preallocation size when
+	doing delayed allocation writeout (default size is 64KiB).
+	Valid values for this option are page size (typically 4KiB)
+	through to 1GiB, inclusive, in power-of-2 increments.
+
+	The default behaviour is for dynamic end-of-file
+	preallocation size, which uses a set of heuristics to
+	optimise the preallocation size based on the current
+	allocation patterns within the file and the access patterns
+	to the file. Specifying a fixed allocsize value turns off
+	the dynamic behaviour.
+
+  attr2
+  noattr2
+	The options enable/disable an "opportunistic" improvement to
+	be made in the way inline extended attributes are stored
+	on-disk.  When the new form is used for the first time when
+	attr2 is selected (either when setting or removing extended
+	attributes) the on-disk superblock feature bit field will be
+	updated to reflect this format being in use.
+
+	The default behaviour is determined by the on-disk feature
+	bit indicating that attr2 behaviour is active. If either
+	mount option it set, then that becomes the new default used
+	by the filesystem.
+
+	CRC enabled filesystems always use the attr2 format, and so
+	will reject the noattr2 mount option if it is set.
+
+  discard
+  nodiscard (*)
+	Enable/disable the issuing of commands to let the block
+	device reclaim space freed by the filesystem.  This is
+	useful for SSD devices, thinly provisioned LUNs and virtual
+	machine images, but may have a performance impact.
+
+	Note: It is currently recommended that you use the fstrim
+	application to discard unused blocks rather than the discard
+	mount option because the performance impact of this option
+	is quite severe.
+
+  grpid/bsdgroups
+  nogrpid/sysvgroups (*)
+	These options define what group ID a newly created file
+	gets.  When grpid is set, it takes the group ID of the
+	directory in which it is created; otherwise it takes the
+	fsgid of the current process, unless the directory has the
+	setgid bit set, in which case it takes the gid from the
+	parent directory, and also gets the setgid bit set if it is
+	a directory itself.
+
+  filestreams
+	Make the data allocator use the filestreams allocation mode
+	across the entire filesystem rather than just on directories
+	configured to use it.
+
+  ikeep
+  noikeep (*)
+	When ikeep is specified, XFS does not delete empty inode
+	clusters and keeps them around on disk.  When noikeep is
+	specified, empty inode clusters are returned to the free
+	space pool.
+
+  inode32
+  inode64 (*)
+	When inode32 is specified, it indicates that XFS limits
+	inode creation to locations which will not result in inode
+	numbers with more than 32 bits of significance.
+
+	When inode64 is specified, it indicates that XFS is allowed
+	to create inodes at any location in the filesystem,
+	including those which will result in inode numbers occupying
+	more than 32 bits of significance. 
+
+	inode32 is provided for backwards compatibility with older
+	systems and applications, since 64 bits inode numbers might
+	cause problems for some applications that cannot handle
+	large inode numbers.  If applications are in use which do
+	not handle inode numbers bigger than 32 bits, the inode32
+	option should be specified.
+
+
+  largeio
+  nolargeio (*)
+	If "nolargeio" is specified, the optimal I/O reported in
+	st_blksize by stat(2) will be as small as possible to allow
+	user applications to avoid inefficient read/modify/write
+	I/O.  This is typically the page size of the machine, as
+	this is the granularity of the page cache.
+
+	If "largeio" specified, a filesystem that was created with a
+	"swidth" specified will return the "swidth" value (in bytes)
+	in st_blksize. If the filesystem does not have a "swidth"
+	specified but does specify an "allocsize" then "allocsize"
+	(in bytes) will be returned instead. Otherwise the behaviour
+	is the same as if "nolargeio" was specified.
+
+  logbufs=value
+	Set the number of in-memory log buffers.  Valid numbers
+	range from 2-8 inclusive.
+
+	The default value is 8 buffers.
+
+	If the memory cost of 8 log buffers is too high on small
+	systems, then it may be reduced at some cost to performance
+	on metadata intensive workloads. The logbsize option below
+	controls the size of each buffer and so is also relevant to
+	this case.
+
+  logbsize=value
+	Set the size of each in-memory log buffer.  The size may be
+	specified in bytes, or in kilobytes with a "k" suffix.
+	Valid sizes for version 1 and version 2 logs are 16384 (16k)
+	and 32768 (32k).  Valid sizes for version 2 logs also
+	include 65536 (64k), 131072 (128k) and 262144 (256k). The
+	logbsize must be an integer multiple of the log
+	stripe unit configured at mkfs time.
+
+	The default value for for version 1 logs is 32768, while the
+	default value for version 2 logs is MAX(32768, log_sunit).
+
+  logdev=device and rtdev=device
+	Use an external log (metadata journal) and/or real-time device.
+	An XFS filesystem has up to three parts: a data section, a log
+	section, and a real-time section.  The real-time section is
+	optional, and the log section can be separate from the data
+	section or contained within it.
+
+  noalign
+	Data allocations will not be aligned at stripe unit
+	boundaries. This is only relevant to filesystems created
+	with non-zero data alignment parameters (sunit, swidth) by
+	mkfs.
+
+  norecovery
+	The filesystem will be mounted without running log recovery.
+	If the filesystem was not cleanly unmounted, it is likely to
+	be inconsistent when mounted in "norecovery" mode.
+	Some files or directories may not be accessible because of this.
+	Filesystems mounted "norecovery" must be mounted read-only or
+	the mount will fail.
+
+  nouuid
+	Don't check for double mounted file systems using the file
+	system uuid.  This is useful to mount LVM snapshot volumes,
+	and often used in combination with "norecovery" for mounting
+	read-only snapshots.
+
+  noquota
+	Forcibly turns off all quota accounting and enforcement
+	within the filesystem.
+
+  uquota/usrquota/uqnoenforce/quota
+	User disk quota accounting enabled, and limits (optionally)
+	enforced.  Refer to xfs_quota(8) for further details.
+
+  gquota/grpquota/gqnoenforce
+	Group disk quota accounting enabled and limits (optionally)
+	enforced.  Refer to xfs_quota(8) for further details.
+
+  pquota/prjquota/pqnoenforce
+	Project disk quota accounting enabled and limits (optionally)
+	enforced.  Refer to xfs_quota(8) for further details.
+
+  sunit=value and swidth=value
+	Used to specify the stripe unit and width for a RAID device
+	or a stripe volume.  "value" must be specified in 512-byte
+	block units. These options are only relevant to filesystems
+	that were created with non-zero data alignment parameters.
+
+	The sunit and swidth parameters specified must be compatible
+	with the existing filesystem alignment characteristics.  In
+	general, that means the only valid changes to sunit are
+	increasing it by a power-of-2 multiple. Valid swidth values
+	are any integer multiple of a valid sunit value.
+
+	Typically the only time these mount options are necessary if
+	after an underlying RAID device has had it's geometry
+	modified, such as adding a new disk to a RAID5 lun and
+	reshaping it.
+
+  swalloc
+	Data allocations will be rounded up to stripe width boundaries
+	when the current end of file is being extended and the file
+	size is larger than the stripe width size.
+
+  wsync
+	When specified, all filesystem namespace operations are
+	executed synchronously. This ensures that when the namespace
+	operation (create, unlink, etc) completes, the change to the
+	namespace is on stable storage. This is useful in HA setups
+	where failover must not result in clients seeing
+	inconsistent namespace presentation during or after a
+	failover event.
+
+
+Deprecated Mount Options
+========================
+
+  Name				Removal Schedule
+  ----				----------------
+
+
+Removed Mount Options
+=====================
+
+  Name				Removed
+  ----				-------
+  delaylog/nodelaylog		v4.0
+  ihashsize			v4.0
+  irixsgid			v4.0
+  osyncisdsync/osyncisosync	v4.0
+  barrier			v4.19
+  nobarrier			v4.19
+
+
+sysctls
+=======
+
+The following sysctls are available for the XFS filesystem:
+
+  fs.xfs.stats_clear		(Min: 0  Default: 0  Max: 1)
+	Setting this to "1" clears accumulated XFS statistics
+	in /proc/fs/xfs/stat.  It then immediately resets to "0".
+
+  fs.xfs.xfssyncd_centisecs	(Min: 100  Default: 3000  Max: 720000)
+	The interval at which the filesystem flushes metadata
+	out to disk and runs internal cache cleanup routines.
+
+  fs.xfs.filestream_centisecs	(Min: 1  Default: 3000  Max: 360000)
+	The interval at which the filesystem ages filestreams cache
+	references and returns timed-out AGs back to the free stream
+	pool.
+
+  fs.xfs.speculative_prealloc_lifetime
+		(Units: seconds   Min: 1  Default: 300  Max: 86400)
+	The interval at which the background scanning for inodes
+	with unused speculative preallocation runs. The scan
+	removes unused preallocation from clean inodes and releases
+	the unused space back to the free pool.
+
+  fs.xfs.error_level		(Min: 0  Default: 3  Max: 11)
+	A volume knob for error reporting when internal errors occur.
+	This will generate detailed messages & backtraces for filesystem
+	shutdowns, for example.  Current threshold values are:
+
+		XFS_ERRLEVEL_OFF:       0
+		XFS_ERRLEVEL_LOW:       1
+		XFS_ERRLEVEL_HIGH:      5
+
+  fs.xfs.panic_mask		(Min: 0  Default: 0  Max: 255)
+	Causes certain error conditions to call BUG(). Value is a bitmask;
+	OR together the tags which represent errors which should cause panics:
+
+		XFS_NO_PTAG                     0
+		XFS_PTAG_IFLUSH                 0x00000001
+		XFS_PTAG_LOGRES                 0x00000002
+		XFS_PTAG_AILDELETE              0x00000004
+		XFS_PTAG_ERROR_REPORT           0x00000008
+		XFS_PTAG_SHUTDOWN_CORRUPT       0x00000010
+		XFS_PTAG_SHUTDOWN_IOERROR       0x00000020
+		XFS_PTAG_SHUTDOWN_LOGERROR      0x00000040
+		XFS_PTAG_FSBLOCK_ZERO           0x00000080
+
+	This option is intended for debugging only.
+
+  fs.xfs.irix_symlink_mode	(Min: 0  Default: 0  Max: 1)
+	Controls whether symlinks are created with mode 0777 (default)
+	or whether their mode is affected by the umask (irix mode).
+
+  fs.xfs.irix_sgid_inherit	(Min: 0  Default: 0  Max: 1)
+	Controls files created in SGID directories.
+	If the group ID of the new file does not match the effective group
+	ID or one of the supplementary group IDs of the parent dir, the
+	ISGID bit is cleared if the irix_sgid_inherit compatibility sysctl
+	is set.
+
+  fs.xfs.inherit_sync		(Min: 0  Default: 1  Max: 1)
+	Setting this to "1" will cause the "sync" flag set
+	by the xfs_io(8) chattr command on a directory to be
+	inherited by files in that directory.
+
+  fs.xfs.inherit_nodump		(Min: 0  Default: 1  Max: 1)
+	Setting this to "1" will cause the "nodump" flag set
+	by the xfs_io(8) chattr command on a directory to be
+	inherited by files in that directory.
+
+  fs.xfs.inherit_noatime	(Min: 0  Default: 1  Max: 1)
+	Setting this to "1" will cause the "noatime" flag set
+	by the xfs_io(8) chattr command on a directory to be
+	inherited by files in that directory.
+
+  fs.xfs.inherit_nosymlinks	(Min: 0  Default: 1  Max: 1)
+	Setting this to "1" will cause the "nosymlinks" flag set
+	by the xfs_io(8) chattr command on a directory to be
+	inherited by files in that directory.
+
+  fs.xfs.inherit_nodefrag	(Min: 0  Default: 1  Max: 1)
+	Setting this to "1" will cause the "nodefrag" flag set
+	by the xfs_io(8) chattr command on a directory to be
+	inherited by files in that directory.
+
+  fs.xfs.rotorstep		(Min: 1  Default: 1  Max: 256)
+	In "inode32" allocation mode, this option determines how many
+	files the allocator attempts to allocate in the same allocation
+	group before moving to the next allocation group.  The intent
+	is to control the rate at which the allocator moves between
+	allocation groups when allocating extents for new files.
+
+Deprecated Sysctls
+==================
+
+None at present.
+
+
+Removed Sysctls
+===============
+
+  Name				Removed
+  ----				-------
+  fs.xfs.xfsbufd_centisec	v4.0
+  fs.xfs.age_buffer_centisecs	v4.0
+
+
+Error handling
+==============
+
+XFS can act differently according to the type of error found during its
+operation. The implementation introduces the following concepts to the error
+handler:
+
+ -failure speed:
+	Defines how fast XFS should propagate an error upwards when a specific
+	error is found during the filesystem operation. It can propagate
+	immediately, after a defined number of retries, after a set time period,
+	or simply retry forever.
+
+ -error classes:
+	Specifies the subsystem the error configuration will apply to, such as
+	metadata IO or memory allocation. Different subsystems will have
+	different error handlers for which behaviour can be configured.
+
+ -error handlers:
+	Defines the behavior for a specific error.
+
+The filesystem behavior during an error can be set via sysfs files. Each
+error handler works independently - the first condition met by an error handler
+for a specific class will cause the error to be propagated rather than reset and
+retried.
+
+The action taken by the filesystem when the error is propagated is context
+dependent - it may cause a shut down in the case of an unrecoverable error,
+it may be reported back to userspace, or it may even be ignored because
+there's nothing useful we can with the error or anyone we can report it to (e.g.
+during unmount).
+
+The configuration files are organized into the following hierarchy for each
+mounted filesystem:
+
+  /sys/fs/xfs/<dev>/error/<class>/<error>/
+
+Where:
+  <dev>
+	The short device name of the mounted filesystem. This is the same device
+	name that shows up in XFS kernel error messages as "XFS(<dev>): ..."
+
+  <class>
+	The subsystem the error configuration belongs to. As of 4.9, the defined
+	classes are:
+
+		- "metadata": applies metadata buffer write IO
+
+  <error>
+	The individual error handler configurations.
+
+
+Each filesystem has "global" error configuration options defined in their top
+level directory:
+
+  /sys/fs/xfs/<dev>/error/
+
+  fail_at_unmount		(Min:  0  Default:  1  Max: 1)
+	Defines the filesystem error behavior at unmount time.
+
+	If set to a value of 1, XFS will override all other error configurations
+	during unmount and replace them with "immediate fail" characteristics.
+	i.e. no retries, no retry timeout. This will always allow unmount to
+	succeed when there are persistent errors present.
+
+	If set to 0, the configured retry behaviour will continue until all
+	retries and/or timeouts have been exhausted. This will delay unmount
+	completion when there are persistent errors, and it may prevent the
+	filesystem from ever unmounting fully in the case of "retry forever"
+	handler configurations.
+
+	Note: there is no guarantee that fail_at_unmount can be set whilst an
+	unmount is in progress. It is possible that the sysfs entries are
+	removed by the unmounting filesystem before a "retry forever" error
+	handler configuration causes unmount to hang, and hence the filesystem
+	must be configured appropriately before unmount begins to prevent
+	unmount hangs.
+
+Each filesystem has specific error class handlers that define the error
+propagation behaviour for specific errors. There is also a "default" error
+handler defined, which defines the behaviour for all errors that don't have
+specific handlers defined. Where multiple retry constraints are configuredi for
+a single error, the first retry configuration that expires will cause the error
+to be propagated. The handler configurations are found in the directory:
+
+  /sys/fs/xfs/<dev>/error/<class>/<error>/
+
+  max_retries			(Min: -1  Default: Varies  Max: INTMAX)
+	Defines the allowed number of retries of a specific error before
+	the filesystem will propagate the error. The retry count for a given
+	error context (e.g. a specific metadata buffer) is reset every time
+	there is a successful completion of the operation.
+
+	Setting the value to "-1" will cause XFS to retry forever for this
+	specific error.
+
+	Setting the value to "0" will cause XFS to fail immediately when the
+	specific error is reported.
+
+	Setting the value to "N" (where 0 < N < Max) will make XFS retry the
+	operation "N" times before propagating the error.
+
+  retry_timeout_seconds		(Min:  -1  Default:  Varies  Max: 1 day)
+	Define the amount of time (in seconds) that the filesystem is
+	allowed to retry its operations when the specific error is
+	found.
+
+	Setting the value to "-1" will allow XFS to retry forever for this
+	specific error.
+
+	Setting the value to "0" will cause XFS to fail immediately when the
+	specific error is reported.
+
+	Setting the value to "N" (where 0 < N < Max) will allow XFS to retry the
+	operation for up to "N" seconds before propagating the error.
+
+Note: The default behaviour for a specific error handler is dependent on both
+the class and error context. For example, the default values for
+"metadata/ENODEV" are "0" rather than "-1" so that this error handler defaults
+to "fail immediately" behaviour. This is done because ENODEV is a fatal,
+unrecoverable error no matter how many times the metadata IO is retried.