Adding upstream version 4.19.249.upstream/4.19.249upstream

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

10 files changed, 2355 insertions, 0 deletions

diff --git a/Documentation/admin-guide/mm/concepts.rst b/Documentation/admin-guide/mm/concepts.rst
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+.. _mm_concepts:
+
+=================
+Concepts overview
+=================
+
+The memory management in Linux is complex system that evolved over the
+years and included more and more functionality to support variety of
+systems from MMU-less microcontrollers to supercomputers. The memory
+management for systems without MMU is called ``nommu`` and it
+definitely deserves a dedicated document, which hopefully will be
+eventually written. Yet, although some of the concepts are the same,
+here we assume that MMU is available and CPU can translate a virtual
+address to a physical address.
+
+.. contents:: :local:
+
+Virtual Memory Primer
+=====================
+
+The physical memory in a computer system is a limited resource and
+even for systems that support memory hotplug there is a hard limit on
+the amount of memory that can be installed. The physical memory is not
+necessary contiguous, it might be accessible as a set of distinct
+address ranges. Besides, different CPU architectures, and even
+different implementations of the same architecture have different view
+how these address ranges defined.
+
+All this makes dealing directly with physical memory quite complex and
+to avoid this complexity a concept of virtual memory was developed.
+
+The virtual memory abstracts the details of physical memory from the
+application software, allows to keep only needed information in the
+physical memory (demand paging) and provides a mechanism for the
+protection and controlled sharing of data between processes.
+
+With virtual memory, each and every memory access uses a virtual
+address. When the CPU decodes the an instruction that reads (or
+writes) from (or to) the system memory, it translates the `virtual`
+address encoded in that instruction to a `physical` address that the
+memory controller can understand.
+
+The physical system memory is divided into page frames, or pages. The
+size of each page is architecture specific. Some architectures allow
+selection of the page size from several supported values; this
+selection is performed at the kernel build time by setting an
+appropriate kernel configuration option.
+
+Each physical memory page can be mapped as one or more virtual
+pages. These mappings are described by page tables that allow
+translation from virtual address used by programs to real address in
+the physical memory. The page tables organized hierarchically.
+
+The tables at the lowest level of the hierarchy contain physical
+addresses of actual pages used by the software. The tables at higher
+levels contain physical addresses of the pages belonging to the lower
+levels. The pointer to the top level page table resides in a
+register. When the CPU performs the address translation, it uses this
+register to access the top level page table. The high bits of the
+virtual address are used to index an entry in the top level page
+table. That entry is then used to access the next level in the
+hierarchy with the next bits of the virtual address as the index to
+that level page table. The lowest bits in the virtual address define
+the offset inside the actual page.
+
+Huge Pages
+==========
+
+The address translation requires several memory accesses and memory
+accesses are slow relatively to CPU speed. To avoid spending precious
+processor cycles on the address translation, CPUs maintain a cache of
+such translations called Translation Lookaside Buffer (or
+TLB). Usually TLB is pretty scarce resource and applications with
+large memory working set will experience performance hit because of
+TLB misses.
+
+Many modern CPU architectures allow mapping of the memory pages
+directly by the higher levels in the page table. For instance, on x86,
+it is possible to map 2M and even 1G pages using entries in the second
+and the third level page tables. In Linux such pages are called
+`huge`. Usage of huge pages significantly reduces pressure on TLB,
+improves TLB hit-rate and thus improves overall system performance.
+
+There are two mechanisms in Linux that enable mapping of the physical
+memory with the huge pages. The first one is `HugeTLB filesystem`, or
+hugetlbfs. It is a pseudo filesystem that uses RAM as its backing
+store. For the files created in this filesystem the data resides in
+the memory and mapped using huge pages. The hugetlbfs is described at
+:ref:`Documentation/admin-guide/mm/hugetlbpage.rst <hugetlbpage>`.
+
+Another, more recent, mechanism that enables use of the huge pages is
+called `Transparent HugePages`, or THP. Unlike the hugetlbfs that
+requires users and/or system administrators to configure what parts of
+the system memory should and can be mapped by the huge pages, THP
+manages such mappings transparently to the user and hence the
+name. See
+:ref:`Documentation/admin-guide/mm/transhuge.rst <admin_guide_transhuge>`
+for more details about THP.
+
+Zones
+=====
+
+Often hardware poses restrictions on how different physical memory
+ranges can be accessed. In some cases, devices cannot perform DMA to
+all the addressable memory. In other cases, the size of the physical
+memory exceeds the maximal addressable size of virtual memory and
+special actions are required to access portions of the memory. Linux
+groups memory pages into `zones` according to their possible
+usage. For example, ZONE_DMA will contain memory that can be used by
+devices for DMA, ZONE_HIGHMEM will contain memory that is not
+permanently mapped into kernel's address space and ZONE_NORMAL will
+contain normally addressed pages.
+
+The actual layout of the memory zones is hardware dependent as not all
+architectures define all zones, and requirements for DMA are different
+for different platforms.
+
+Nodes
+=====
+
+Many multi-processor machines are NUMA - Non-Uniform Memory Access -
+systems. In such systems the memory is arranged into banks that have
+different access latency depending on the "distance" from the
+processor. Each bank is referred as `node` and for each node Linux
+constructs an independent memory management subsystem. A node has it's
+own set of zones, lists of free and used pages and various statistics
+counters. You can find more details about NUMA in
+:ref:`Documentation/vm/numa.rst <numa>` and in
+:ref:`Documentation/admin-guide/mm/numa_memory_policy.rst <numa_memory_policy>`.
+
+Page cache
+==========
+
+The physical memory is volatile and the common case for getting data
+into the memory is to read it from files. Whenever a file is read, the
+data is put into the `page cache` to avoid expensive disk access on
+the subsequent reads. Similarly, when one writes to a file, the data
+is placed in the page cache and eventually gets into the backing
+storage device. The written pages are marked as `dirty` and when Linux
+decides to reuse them for other purposes, it makes sure to synchronize
+the file contents on the device with the updated data.
+
+Anonymous Memory
+================
+
+The `anonymous memory` or `anonymous mappings` represent memory that
+is not backed by a filesystem. Such mappings are implicitly created
+for program's stack and heap or by explicit calls to mmap(2) system
+call. Usually, the anonymous mappings only define virtual memory areas
+that the program is allowed to access. The read accesses will result
+in creation of a page table entry that references a special physical
+page filled with zeroes. When the program performs a write, regular
+physical page will be allocated to hold the written data. The page
+will be marked dirty and if the kernel will decide to repurpose it,
+the dirty page will be swapped out.
+
+Reclaim
+=======
+
+Throughout the system lifetime, a physical page can be used for storing
+different types of data. It can be kernel internal data structures,
+DMA'able buffers for device drivers use, data read from a filesystem,
+memory allocated by user space processes etc.
+
+Depending on the page usage it is treated differently by the Linux
+memory management. The pages that can be freed at any time, either
+because they cache the data available elsewhere, for instance, on a
+hard disk, or because they can be swapped out, again, to the hard
+disk, are called `reclaimable`. The most notable categories of the
+reclaimable pages are page cache and anonymous memory.
+
+In most cases, the pages holding internal kernel data and used as DMA
+buffers cannot be repurposed, and they remain pinned until freed by
+their user. Such pages are called `unreclaimable`. However, in certain
+circumstances, even pages occupied with kernel data structures can be
+reclaimed. For instance, in-memory caches of filesystem metadata can
+be re-read from the storage device and therefore it is possible to
+discard them from the main memory when system is under memory
+pressure.
+
+The process of freeing the reclaimable physical memory pages and
+repurposing them is called (surprise!) `reclaim`. Linux can reclaim
+pages either asynchronously or synchronously, depending on the state
+of the system. When system is not loaded, most of the memory is free
+and allocation request will be satisfied immediately from the free
+pages supply. As the load increases, the amount of the free pages goes
+down and when it reaches a certain threshold (high watermark), an
+allocation request will awaken the ``kswapd`` daemon. It will
+asynchronously scan memory pages and either just free them if the data
+they contain is available elsewhere, or evict to the backing storage
+device (remember those dirty pages?). As memory usage increases even
+more and reaches another threshold - min watermark - an allocation
+will trigger the `direct reclaim`. In this case allocation is stalled
+until enough memory pages are reclaimed to satisfy the request.
+
+Compaction
+==========
+
+As the system runs, tasks allocate and free the memory and it becomes
+fragmented. Although with virtual memory it is possible to present
+scattered physical pages as virtually contiguous range, sometimes it is
+necessary to allocate large physically contiguous memory areas. Such
+need may arise, for instance, when a device driver requires large
+buffer for DMA, or when THP allocates a huge page. Memory `compaction`
+addresses the fragmentation issue. This mechanism moves occupied pages
+from the lower part of a memory zone to free pages in the upper part
+of the zone. When a compaction scan is finished free pages are grouped
+together at the beginning of the zone and allocations of large
+physically contiguous areas become possible.
+
+Like reclaim, the compaction may happen asynchronously in ``kcompactd``
+daemon or synchronously as a result of memory allocation request.
+
+OOM killer
+==========
+
+It may happen, that on a loaded machine memory will be exhausted. When
+the kernel detects that the system runs out of memory (OOM) it invokes
+`OOM killer`. Its mission is simple: all it has to do is to select a
+task to sacrifice for the sake of the overall system health. The
+selected task is killed in a hope that after it exits enough memory
+will be freed to continue normal operation.
diff --git a/Documentation/admin-guide/mm/hugetlbpage.rst b/Documentation/admin-guide/mm/hugetlbpage.rst
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index 000000000..1cc0bc78d
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+++ b/Documentation/admin-guide/mm/hugetlbpage.rst
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+.. _hugetlbpage:
+
+=============
+HugeTLB Pages
+=============
+
+Overview
+========
+
+The intent of this file is to give a brief summary of hugetlbpage support in
+the Linux kernel.  This support is built on top of multiple page size support
+that is provided by most modern architectures.  For example, x86 CPUs normally
+support 4K and 2M (1G if architecturally supported) page sizes, ia64
+architecture supports multiple page sizes 4K, 8K, 64K, 256K, 1M, 4M, 16M,
+256M and ppc64 supports 4K and 16M.  A TLB is a cache of virtual-to-physical
+translations.  Typically this is a very scarce resource on processor.
+Operating systems try to make best use of limited number of TLB resources.
+This optimization is more critical now as bigger and bigger physical memories
+(several GBs) are more readily available.
+
+Users can use the huge page support in Linux kernel by either using the mmap
+system call or standard SYSV shared memory system calls (shmget, shmat).
+
+First the Linux kernel needs to be built with the CONFIG_HUGETLBFS
+(present under "File systems") and CONFIG_HUGETLB_PAGE (selected
+automatically when CONFIG_HUGETLBFS is selected) configuration
+options.
+
+The ``/proc/meminfo`` file provides information about the total number of
+persistent hugetlb pages in the kernel's huge page pool.  It also displays
+default huge page size and information about the number of free, reserved
+and surplus huge pages in the pool of huge pages of default size.
+The huge page size is needed for generating the proper alignment and
+size of the arguments to system calls that map huge page regions.
+
+The output of ``cat /proc/meminfo`` will include lines like::
+
+	HugePages_Total: uuu
+	HugePages_Free:  vvv
+	HugePages_Rsvd:  www
+	HugePages_Surp:  xxx
+	Hugepagesize:    yyy kB
+	Hugetlb:         zzz kB
+
+where:
+
+HugePages_Total
+	is the size of the pool of huge pages.
+HugePages_Free
+	is the number of huge pages in the pool that are not yet
+        allocated.
+HugePages_Rsvd
+	is short for "reserved," and is the number of huge pages for
+        which a commitment to allocate from the pool has been made,
+        but no allocation has yet been made.  Reserved huge pages
+        guarantee that an application will be able to allocate a
+        huge page from the pool of huge pages at fault time.
+HugePages_Surp
+	is short for "surplus," and is the number of huge pages in
+        the pool above the value in ``/proc/sys/vm/nr_hugepages``. The
+        maximum number of surplus huge pages is controlled by
+        ``/proc/sys/vm/nr_overcommit_hugepages``.
+Hugepagesize
+	is the default hugepage size (in Kb).
+Hugetlb
+        is the total amount of memory (in kB), consumed by huge
+        pages of all sizes.
+        If huge pages of different sizes are in use, this number
+        will exceed HugePages_Total \* Hugepagesize. To get more
+        detailed information, please, refer to
+        ``/sys/kernel/mm/hugepages`` (described below).
+
+
+``/proc/filesystems`` should also show a filesystem of type "hugetlbfs"
+configured in the kernel.
+
+``/proc/sys/vm/nr_hugepages`` indicates the current number of "persistent" huge
+pages in the kernel's huge page pool.  "Persistent" huge pages will be
+returned to the huge page pool when freed by a task.  A user with root
+privileges can dynamically allocate more or free some persistent huge pages
+by increasing or decreasing the value of ``nr_hugepages``.
+
+Pages that are used as huge pages are reserved inside the kernel and cannot
+be used for other purposes.  Huge pages cannot be swapped out under
+memory pressure.
+
+Once a number of huge pages have been pre-allocated to the kernel huge page
+pool, a user with appropriate privilege can use either the mmap system call
+or shared memory system calls to use the huge pages.  See the discussion of
+:ref:`Using Huge Pages <using_huge_pages>`, below.
+
+The administrator can allocate persistent huge pages on the kernel boot
+command line by specifying the "hugepages=N" parameter, where 'N' = the
+number of huge pages requested.  This is the most reliable method of
+allocating huge pages as memory has not yet become fragmented.
+
+Some platforms support multiple huge page sizes.  To allocate huge pages
+of a specific size, one must precede the huge pages boot command parameters
+with a huge page size selection parameter "hugepagesz=<size>".  <size> must
+be specified in bytes with optional scale suffix [kKmMgG].  The default huge
+page size may be selected with the "default_hugepagesz=<size>" boot parameter.
+
+When multiple huge page sizes are supported, ``/proc/sys/vm/nr_hugepages``
+indicates the current number of pre-allocated huge pages of the default size.
+Thus, one can use the following command to dynamically allocate/deallocate
+default sized persistent huge pages::
+
+	echo 20 > /proc/sys/vm/nr_hugepages
+
+This command will try to adjust the number of default sized huge pages in the
+huge page pool to 20, allocating or freeing huge pages, as required.
+
+On a NUMA platform, the kernel will attempt to distribute the huge page pool
+over all the set of allowed nodes specified by the NUMA memory policy of the
+task that modifies ``nr_hugepages``. The default for the allowed nodes--when the
+task has default memory policy--is all on-line nodes with memory.  Allowed
+nodes with insufficient available, contiguous memory for a huge page will be
+silently skipped when allocating persistent huge pages.  See the
+:ref:`discussion below <mem_policy_and_hp_alloc>`
+of the interaction of task memory policy, cpusets and per node attributes
+with the allocation and freeing of persistent huge pages.
+
+The success or failure of huge page allocation depends on the amount of
+physically contiguous memory that is present in system at the time of the
+allocation attempt.  If the kernel is unable to allocate huge pages from
+some nodes in a NUMA system, it will attempt to make up the difference by
+allocating extra pages on other nodes with sufficient available contiguous
+memory, if any.
+
+System administrators may want to put this command in one of the local rc
+init files.  This will enable the kernel to allocate huge pages early in
+the boot process when the possibility of getting physical contiguous pages
+is still very high.  Administrators can verify the number of huge pages
+actually allocated by checking the sysctl or meminfo.  To check the per node
+distribution of huge pages in a NUMA system, use::
+
+	cat /sys/devices/system/node/node*/meminfo | fgrep Huge
+
+``/proc/sys/vm/nr_overcommit_hugepages`` specifies how large the pool of
+huge pages can grow, if more huge pages than ``/proc/sys/vm/nr_hugepages`` are
+requested by applications.  Writing any non-zero value into this file
+indicates that the hugetlb subsystem is allowed to try to obtain that
+number of "surplus" huge pages from the kernel's normal page pool, when the
+persistent huge page pool is exhausted. As these surplus huge pages become
+unused, they are freed back to the kernel's normal page pool.
+
+When increasing the huge page pool size via ``nr_hugepages``, any existing
+surplus pages will first be promoted to persistent huge pages.  Then, additional
+huge pages will be allocated, if necessary and if possible, to fulfill
+the new persistent huge page pool size.
+
+The administrator may shrink the pool of persistent huge pages for
+the default huge page size by setting the ``nr_hugepages`` sysctl to a
+smaller value.  The kernel will attempt to balance the freeing of huge pages
+across all nodes in the memory policy of the task modifying ``nr_hugepages``.
+Any free huge pages on the selected nodes will be freed back to the kernel's
+normal page pool.
+
+Caveat: Shrinking the persistent huge page pool via ``nr_hugepages`` such that
+it becomes less than the number of huge pages in use will convert the balance
+of the in-use huge pages to surplus huge pages.  This will occur even if
+the number of surplus pages would exceed the overcommit value.  As long as
+this condition holds--that is, until ``nr_hugepages+nr_overcommit_hugepages`` is
+increased sufficiently, or the surplus huge pages go out of use and are freed--
+no more surplus huge pages will be allowed to be allocated.
+
+With support for multiple huge page pools at run-time available, much of
+the huge page userspace interface in ``/proc/sys/vm`` has been duplicated in
+sysfs.
+The ``/proc`` interfaces discussed above have been retained for backwards
+compatibility. The root huge page control directory in sysfs is::
+
+	/sys/kernel/mm/hugepages
+
+For each huge page size supported by the running kernel, a subdirectory
+will exist, of the form::
+
+	hugepages-${size}kB
+
+Inside each of these directories, the same set of files will exist::
+
+	nr_hugepages
+	nr_hugepages_mempolicy
+	nr_overcommit_hugepages
+	free_hugepages
+	resv_hugepages
+	surplus_hugepages
+
+which function as described above for the default huge page-sized case.
+
+.. _mem_policy_and_hp_alloc:
+
+Interaction of Task Memory Policy with Huge Page Allocation/Freeing
+===================================================================
+
+Whether huge pages are allocated and freed via the ``/proc`` interface or
+the ``/sysfs`` interface using the ``nr_hugepages_mempolicy`` attribute, the
+NUMA nodes from which huge pages are allocated or freed are controlled by the
+NUMA memory policy of the task that modifies the ``nr_hugepages_mempolicy``
+sysctl or attribute.  When the ``nr_hugepages`` attribute is used, mempolicy
+is ignored.
+
+The recommended method to allocate or free huge pages to/from the kernel
+huge page pool, using the ``nr_hugepages`` example above, is::
+
+    numactl --interleave <node-list> echo 20 \
+				>/proc/sys/vm/nr_hugepages_mempolicy
+
+or, more succinctly::
+
+    numactl -m <node-list> echo 20 >/proc/sys/vm/nr_hugepages_mempolicy
+
+This will allocate or free ``abs(20 - nr_hugepages)`` to or from the nodes
+specified in <node-list>, depending on whether number of persistent huge pages
+is initially less than or greater than 20, respectively.  No huge pages will be
+allocated nor freed on any node not included in the specified <node-list>.
+
+When adjusting the persistent hugepage count via ``nr_hugepages_mempolicy``, any
+memory policy mode--bind, preferred, local or interleave--may be used.  The
+resulting effect on persistent huge page allocation is as follows:
+
+#. Regardless of mempolicy mode [see
+   :ref:`Documentation/admin-guide/mm/numa_memory_policy.rst <numa_memory_policy>`],
+   persistent huge pages will be distributed across the node or nodes
+   specified in the mempolicy as if "interleave" had been specified.
+   However, if a node in the policy does not contain sufficient contiguous
+   memory for a huge page, the allocation will not "fallback" to the nearest
+   neighbor node with sufficient contiguous memory.  To do this would cause
+   undesirable imbalance in the distribution of the huge page pool, or
+   possibly, allocation of persistent huge pages on nodes not allowed by
+   the task's memory policy.
+
+#. One or more nodes may be specified with the bind or interleave policy.
+   If more than one node is specified with the preferred policy, only the
+   lowest numeric id will be used.  Local policy will select the node where
+   the task is running at the time the nodes_allowed mask is constructed.
+   For local policy to be deterministic, the task must be bound to a cpu or
+   cpus in a single node.  Otherwise, the task could be migrated to some
+   other node at any time after launch and the resulting node will be
+   indeterminate.  Thus, local policy is not very useful for this purpose.
+   Any of the other mempolicy modes may be used to specify a single node.
+
+#. The nodes allowed mask will be derived from any non-default task mempolicy,
+   whether this policy was set explicitly by the task itself or one of its
+   ancestors, such as numactl.  This means that if the task is invoked from a
+   shell with non-default policy, that policy will be used.  One can specify a
+   node list of "all" with numactl --interleave or --membind [-m] to achieve
+   interleaving over all nodes in the system or cpuset.
+
+#. Any task mempolicy specified--e.g., using numactl--will be constrained by
+   the resource limits of any cpuset in which the task runs.  Thus, there will
+   be no way for a task with non-default policy running in a cpuset with a
+   subset of the system nodes to allocate huge pages outside the cpuset
+   without first moving to a cpuset that contains all of the desired nodes.
+
+#. Boot-time huge page allocation attempts to distribute the requested number
+   of huge pages over all on-lines nodes with memory.
+
+Per Node Hugepages Attributes
+=============================
+
+A subset of the contents of the root huge page control directory in sysfs,
+described above, will be replicated under each the system device of each
+NUMA node with memory in::
+
+	/sys/devices/system/node/node[0-9]*/hugepages/
+
+Under this directory, the subdirectory for each supported huge page size
+contains the following attribute files::
+
+	nr_hugepages
+	free_hugepages
+	surplus_hugepages
+
+The free\_' and surplus\_' attribute files are read-only.  They return the number
+of free and surplus [overcommitted] huge pages, respectively, on the parent
+node.
+
+The ``nr_hugepages`` attribute returns the total number of huge pages on the
+specified node.  When this attribute is written, the number of persistent huge
+pages on the parent node will be adjusted to the specified value, if sufficient
+resources exist, regardless of the task's mempolicy or cpuset constraints.
+
+Note that the number of overcommit and reserve pages remain global quantities,
+as we don't know until fault time, when the faulting task's mempolicy is
+applied, from which node the huge page allocation will be attempted.
+
+.. _using_huge_pages:
+
+Using Huge Pages
+================
+
+If the user applications are going to request huge pages using mmap system
+call, then it is required that system administrator mount a file system of
+type hugetlbfs::
+
+  mount -t hugetlbfs \
+	-o uid=<value>,gid=<value>,mode=<value>,pagesize=<value>,size=<value>,\
+	min_size=<value>,nr_inodes=<value> none /mnt/huge
+
+This command mounts a (pseudo) filesystem of type hugetlbfs on the directory
+``/mnt/huge``.  Any file created on ``/mnt/huge`` uses huge pages.
+
+The ``uid`` and ``gid`` options sets the owner and group of the root of the
+file system.  By default the ``uid`` and ``gid`` of the current process
+are taken.
+
+The ``mode`` option sets the mode of root of file system to value & 01777.
+This value is given in octal. By default the value 0755 is picked.
+
+If the platform supports multiple huge page sizes, the ``pagesize`` option can
+be used to specify the huge page size and associated pool. ``pagesize``
+is specified in bytes. If ``pagesize`` is not specified the platform's
+default huge page size and associated pool will be used.
+
+The ``size`` option sets the maximum value of memory (huge pages) allowed
+for that filesystem (``/mnt/huge``). The ``size`` option can be specified
+in bytes, or as a percentage of the specified huge page pool (``nr_hugepages``).
+The size is rounded down to HPAGE_SIZE boundary.
+
+The ``min_size`` option sets the minimum value of memory (huge pages) allowed
+for the filesystem. ``min_size`` can be specified in the same way as ``size``,
+either bytes or a percentage of the huge page pool.
+At mount time, the number of huge pages specified by ``min_size`` are reserved
+for use by the filesystem.
+If there are not enough free huge pages available, the mount will fail.
+As huge pages are allocated to the filesystem and freed, the reserve count
+is adjusted so that the sum of allocated and reserved huge pages is always
+at least ``min_size``.
+
+The option ``nr_inodes`` sets the maximum number of inodes that ``/mnt/huge``
+can use.
+
+If the ``size``, ``min_size`` or ``nr_inodes`` option is not provided on
+command line then no limits are set.
+
+For ``pagesize``, ``size``, ``min_size`` and ``nr_inodes`` options, you can
+use [G|g]/[M|m]/[K|k] to represent giga/mega/kilo.
+For example, size=2K has the same meaning as size=2048.
+
+While read system calls are supported on files that reside on hugetlb
+file systems, write system calls are not.
+
+Regular chown, chgrp, and chmod commands (with right permissions) could be
+used to change the file attributes on hugetlbfs.
+
+Also, it is important to note that no such mount command is required if
+applications are going to use only shmat/shmget system calls or mmap with
+MAP_HUGETLB.  For an example of how to use mmap with MAP_HUGETLB see
+:ref:`map_hugetlb <map_hugetlb>` below.
+
+Users who wish to use hugetlb memory via shared memory segment should be
+members of a supplementary group and system admin needs to configure that gid
+into ``/proc/sys/vm/hugetlb_shm_group``.  It is possible for same or different
+applications to use any combination of mmaps and shm* calls, though the mount of
+filesystem will be required for using mmap calls without MAP_HUGETLB.
+
+Syscalls that operate on memory backed by hugetlb pages only have their lengths
+aligned to the native page size of the processor; they will normally fail with
+errno set to EINVAL or exclude hugetlb pages that extend beyond the length if
+not hugepage aligned.  For example, munmap(2) will fail if memory is backed by
+a hugetlb page and the length is smaller than the hugepage size.
+
+
+Examples
+========
+
+.. _map_hugetlb:
+
+``map_hugetlb``
+	see tools/testing/selftests/vm/map_hugetlb.c
+
+``hugepage-shm``
+	see tools/testing/selftests/vm/hugepage-shm.c
+
+``hugepage-mmap``
+	see tools/testing/selftests/vm/hugepage-mmap.c
+
+The `libhugetlbfs`_  library provides a wide range of userspace tools
+to help with huge page usability, environment setup, and control.
+
+.. _libhugetlbfs: https://github.com/libhugetlbfs/libhugetlbfs
diff --git a/Documentation/admin-guide/mm/idle_page_tracking.rst b/Documentation/admin-guide/mm/idle_page_tracking.rst
new file mode 100644
index 000000000..df9394fb3
--- /dev/null
+++ b/Documentation/admin-guide/mm/idle_page_tracking.rst
@@ -0,0 +1,121 @@
+.. _idle_page_tracking:
+
+==================
+Idle Page Tracking
+==================
+
+Motivation
+==========
+
+The idle page tracking feature allows to track which memory pages are being
+accessed by a workload and which are idle. This information can be useful for
+estimating the workload's working set size, which, in turn, can be taken into
+account when configuring the workload parameters, setting memory cgroup limits,
+or deciding where to place the workload within a compute cluster.
+
+It is enabled by CONFIG_IDLE_PAGE_TRACKING=y.
+
+.. _user_api:
+
+User API
+========
+
+The idle page tracking API is located at ``/sys/kernel/mm/page_idle``.
+Currently, it consists of the only read-write file,
+``/sys/kernel/mm/page_idle/bitmap``.
+
+The file implements a bitmap where each bit corresponds to a memory page. The
+bitmap is represented by an array of 8-byte integers, and the page at PFN #i is
+mapped to bit #i%64 of array element #i/64, byte order is native. When a bit is
+set, the corresponding page is idle.
+
+A page is considered idle if it has not been accessed since it was marked idle
+(for more details on what "accessed" actually means see the :ref:`Implementation
+Details <impl_details>` section).
+To mark a page idle one has to set the bit corresponding to
+the page by writing to the file. A value written to the file is OR-ed with the
+current bitmap value.
+
+Only accesses to user memory pages are tracked. These are pages mapped to a
+process address space, page cache and buffer pages, swap cache pages. For other
+page types (e.g. SLAB pages) an attempt to mark a page idle is silently ignored,
+and hence such pages are never reported idle.
+
+For huge pages the idle flag is set only on the head page, so one has to read
+``/proc/kpageflags`` in order to correctly count idle huge pages.
+
+Reading from or writing to ``/sys/kernel/mm/page_idle/bitmap`` will return
+-EINVAL if you are not starting the read/write on an 8-byte boundary, or
+if the size of the read/write is not a multiple of 8 bytes. Writing to
+this file beyond max PFN will return -ENXIO.
+
+That said, in order to estimate the amount of pages that are not used by a
+workload one should:
+
+ 1. Mark all the workload's pages as idle by setting corresponding bits in
+    ``/sys/kernel/mm/page_idle/bitmap``. The pages can be found by reading
+    ``/proc/pid/pagemap`` if the workload is represented by a process, or by
+    filtering out alien pages using ``/proc/kpagecgroup`` in case the workload
+    is placed in a memory cgroup.
+
+ 2. Wait until the workload accesses its working set.
+
+ 3. Read ``/sys/kernel/mm/page_idle/bitmap`` and count the number of bits set.
+    If one wants to ignore certain types of pages, e.g. mlocked pages since they
+    are not reclaimable, he or she can filter them out using
+    ``/proc/kpageflags``.
+
+The page-types tool in the tools/vm directory can be used to assist in this.
+If the tool is run initially with the appropriate option, it will mark all the
+queried pages as idle.  Subsequent runs of the tool can then show which pages have
+their idle flag cleared in the interim.
+
+See :ref:`Documentation/admin-guide/mm/pagemap.rst <pagemap>` for more
+information about ``/proc/pid/pagemap``, ``/proc/kpageflags``, and
+``/proc/kpagecgroup``.
+
+.. _impl_details:
+
+Implementation Details
+======================
+
+The kernel internally keeps track of accesses to user memory pages in order to
+reclaim unreferenced pages first on memory shortage conditions. A page is
+considered referenced if it has been recently accessed via a process address
+space, in which case one or more PTEs it is mapped to will have the Accessed bit
+set, or marked accessed explicitly by the kernel (see mark_page_accessed()). The
+latter happens when:
+
+ - a userspace process reads or writes a page using a system call (e.g. read(2)
+   or write(2))
+
+ - a page that is used for storing filesystem buffers is read or written,
+   because a process needs filesystem metadata stored in it (e.g. lists a
+   directory tree)
+
+ - a page is accessed by a device driver using get_user_pages()
+
+When a dirty page is written to swap or disk as a result of memory reclaim or
+exceeding the dirty memory limit, it is not marked referenced.
+
+The idle memory tracking feature adds a new page flag, the Idle flag. This flag
+is set manually, by writing to ``/sys/kernel/mm/page_idle/bitmap`` (see the
+:ref:`User API <user_api>`
+section), and cleared automatically whenever a page is referenced as defined
+above.
+
+When a page is marked idle, the Accessed bit must be cleared in all PTEs it is
+mapped to, otherwise we will not be able to detect accesses to the page coming
+from a process address space. To avoid interference with the reclaimer, which,
+as noted above, uses the Accessed bit to promote actively referenced pages, one
+more page flag is introduced, the Young flag. When the PTE Accessed bit is
+cleared as a result of setting or updating a page's Idle flag, the Young flag
+is set on the page. The reclaimer treats the Young flag as an extra PTE
+Accessed bit and therefore will consider such a page as referenced.
+
+Since the idle memory tracking feature is based on the memory reclaimer logic,
+it only works with pages that are on an LRU list, other pages are silently
+ignored. That means it will ignore a user memory page if it is isolated, but
+since there are usually not many of them, it should not affect the overall
+result noticeably. In order not to stall scanning of the idle page bitmap,
+locked pages may be skipped too.
diff --git a/Documentation/admin-guide/mm/index.rst b/Documentation/admin-guide/mm/index.rst
new file mode 100644
index 000000000..ceead68c2
--- /dev/null
+++ b/Documentation/admin-guide/mm/index.rst
@@ -0,0 +1,36 @@
+=================
+Memory Management
+=================
+
+Linux memory management subsystem is responsible, as the name implies,
+for managing the memory in the system. This includes implemnetation of
+virtual memory and demand paging, memory allocation both for kernel
+internal structures and user space programms, mapping of files into
+processes address space and many other cool things.
+
+Linux memory management is a complex system with many configurable
+settings. Most of these settings are available via ``/proc``
+filesystem and can be quired and adjusted using ``sysctl``. These APIs
+are described in Documentation/sysctl/vm.txt and in `man 5 proc`_.
+
+.. _man 5 proc: http://man7.org/linux/man-pages/man5/proc.5.html
+
+Linux memory management has its own jargon and if you are not yet
+familiar with it, consider reading
+:ref:`Documentation/admin-guide/mm/concepts.rst <mm_concepts>`.
+
+Here we document in detail how to interact with various mechanisms in
+the Linux memory management.
+
+.. toctree::
+   :maxdepth: 1
+
+   concepts
+   hugetlbpage
+   idle_page_tracking
+   ksm
+   numa_memory_policy
+   pagemap
+   soft-dirty
+   transhuge
+   userfaultfd
diff --git a/Documentation/admin-guide/mm/ksm.rst b/Documentation/admin-guide/mm/ksm.rst
new file mode 100644
index 000000000..930378663
--- /dev/null
+++ b/Documentation/admin-guide/mm/ksm.rst
@@ -0,0 +1,189 @@
+.. _admin_guide_ksm:
+
+=======================
+Kernel Samepage Merging
+=======================
+
+Overview
+========
+
+KSM is a memory-saving de-duplication feature, enabled by CONFIG_KSM=y,
+added to the Linux kernel in 2.6.32.  See ``mm/ksm.c`` for its implementation,
+and http://lwn.net/Articles/306704/ and http://lwn.net/Articles/330589/
+
+KSM was originally developed for use with KVM (where it was known as
+Kernel Shared Memory), to fit more virtual machines into physical memory,
+by sharing the data common between them.  But it can be useful to any
+application which generates many instances of the same data.
+
+The KSM daemon ksmd periodically scans those areas of user memory
+which have been registered with it, looking for pages of identical
+content which can be replaced by a single write-protected page (which
+is automatically copied if a process later wants to update its
+content). The amount of pages that KSM daemon scans in a single pass
+and the time between the passes are configured using :ref:`sysfs
+intraface <ksm_sysfs>`
+
+KSM only merges anonymous (private) pages, never pagecache (file) pages.
+KSM's merged pages were originally locked into kernel memory, but can now
+be swapped out just like other user pages (but sharing is broken when they
+are swapped back in: ksmd must rediscover their identity and merge again).
+
+Controlling KSM with madvise
+============================
+
+KSM only operates on those areas of address space which an application
+has advised to be likely candidates for merging, by using the madvise(2)
+system call::
+
+	int madvise(addr, length, MADV_MERGEABLE)
+
+The app may call
+
+::
+
+	int madvise(addr, length, MADV_UNMERGEABLE)
+
+to cancel that advice and restore unshared pages: whereupon KSM
+unmerges whatever it merged in that range.  Note: this unmerging call
+may suddenly require more memory than is available - possibly failing
+with EAGAIN, but more probably arousing the Out-Of-Memory killer.
+
+If KSM is not configured into the running kernel, madvise MADV_MERGEABLE
+and MADV_UNMERGEABLE simply fail with EINVAL.  If the running kernel was
+built with CONFIG_KSM=y, those calls will normally succeed: even if the
+the KSM daemon is not currently running, MADV_MERGEABLE still registers
+the range for whenever the KSM daemon is started; even if the range
+cannot contain any pages which KSM could actually merge; even if
+MADV_UNMERGEABLE is applied to a range which was never MADV_MERGEABLE.
+
+If a region of memory must be split into at least one new MADV_MERGEABLE
+or MADV_UNMERGEABLE region, the madvise may return ENOMEM if the process
+will exceed ``vm.max_map_count`` (see Documentation/sysctl/vm.txt).
+
+Like other madvise calls, they are intended for use on mapped areas of
+the user address space: they will report ENOMEM if the specified range
+includes unmapped gaps (though working on the intervening mapped areas),
+and might fail with EAGAIN if not enough memory for internal structures.
+
+Applications should be considerate in their use of MADV_MERGEABLE,
+restricting its use to areas likely to benefit.  KSM's scans may use a lot
+of processing power: some installations will disable KSM for that reason.
+
+.. _ksm_sysfs:
+
+KSM daemon sysfs interface
+==========================
+
+The KSM daemon is controlled by sysfs files in ``/sys/kernel/mm/ksm/``,
+readable by all but writable only by root:
+
+pages_to_scan
+        how many pages to scan before ksmd goes to sleep
+        e.g. ``echo 100 > /sys/kernel/mm/ksm/pages_to_scan``.
+
+        Default: 100 (chosen for demonstration purposes)
+
+sleep_millisecs
+        how many milliseconds ksmd should sleep before next scan
+        e.g. ``echo 20 > /sys/kernel/mm/ksm/sleep_millisecs``
+
+        Default: 20 (chosen for demonstration purposes)
+
+merge_across_nodes
+        specifies if pages from different NUMA nodes can be merged.
+        When set to 0, ksm merges only pages which physically reside
+        in the memory area of same NUMA node. That brings lower
+        latency to access of shared pages. Systems with more nodes, at
+        significant NUMA distances, are likely to benefit from the
+        lower latency of setting 0. Smaller systems, which need to
+        minimize memory usage, are likely to benefit from the greater
+        sharing of setting 1 (default). You may wish to compare how
+        your system performs under each setting, before deciding on
+        which to use. ``merge_across_nodes`` setting can be changed only
+        when there are no ksm shared pages in the system: set run 2 to
+        unmerge pages first, then to 1 after changing
+        ``merge_across_nodes``, to remerge according to the new setting.
+
+        Default: 1 (merging across nodes as in earlier releases)
+
+run
+        * set to 0 to stop ksmd from running but keep merged pages,
+        * set to 1 to run ksmd e.g. ``echo 1 > /sys/kernel/mm/ksm/run``,
+        * set to 2 to stop ksmd and unmerge all pages currently merged, but
+	  leave mergeable areas registered for next run.
+
+        Default: 0 (must be changed to 1 to activate KSM, except if
+        CONFIG_SYSFS is disabled)
+
+use_zero_pages
+        specifies whether empty pages (i.e. allocated pages that only
+        contain zeroes) should be treated specially.  When set to 1,
+        empty pages are merged with the kernel zero page(s) instead of
+        with each other as it would happen normally. This can improve
+        the performance on architectures with coloured zero pages,
+        depending on the workload. Care should be taken when enabling
+        this setting, as it can potentially degrade the performance of
+        KSM for some workloads, for example if the checksums of pages
+        candidate for merging match the checksum of an empty
+        page. This setting can be changed at any time, it is only
+        effective for pages merged after the change.
+
+        Default: 0 (normal KSM behaviour as in earlier releases)
+
+max_page_sharing
+        Maximum sharing allowed for each KSM page. This enforces a
+        deduplication limit to avoid high latency for virtual memory
+        operations that involve traversal of the virtual mappings that
+        share the KSM page. The minimum value is 2 as a newly created
+        KSM page will have at least two sharers. The higher this value
+        the faster KSM will merge the memory and the higher the
+        deduplication factor will be, but the slower the worst case
+        virtual mappings traversal could be for any given KSM
+        page. Slowing down this traversal means there will be higher
+        latency for certain virtual memory operations happening during
+        swapping, compaction, NUMA balancing and page migration, in
+        turn decreasing responsiveness for the caller of those virtual
+        memory operations. The scheduler latency of other tasks not
+        involved with the VM operations doing the virtual mappings
+        traversal is not affected by this parameter as these
+        traversals are always schedule friendly themselves.
+
+stable_node_chains_prune_millisecs
+        specifies how frequently KSM checks the metadata of the pages
+        that hit the deduplication limit for stale information.
+        Smaller milllisecs values will free up the KSM metadata with
+        lower latency, but they will make ksmd use more CPU during the
+        scan. It's a noop if not a single KSM page hit the
+        ``max_page_sharing`` yet.
+
+The effectiveness of KSM and MADV_MERGEABLE is shown in ``/sys/kernel/mm/ksm/``:
+
+pages_shared
+        how many shared pages are being used
+pages_sharing
+        how many more sites are sharing them i.e. how much saved
+pages_unshared
+        how many pages unique but repeatedly checked for merging
+pages_volatile
+        how many pages changing too fast to be placed in a tree
+full_scans
+        how many times all mergeable areas have been scanned
+stable_node_chains
+        the number of KSM pages that hit the ``max_page_sharing`` limit
+stable_node_dups
+        number of duplicated KSM pages
+
+A high ratio of ``pages_sharing`` to ``pages_shared`` indicates good
+sharing, but a high ratio of ``pages_unshared`` to ``pages_sharing``
+indicates wasted effort.  ``pages_volatile`` embraces several
+different kinds of activity, but a high proportion there would also
+indicate poor use of madvise MADV_MERGEABLE.
+
+The maximum possible ``pages_sharing/pages_shared`` ratio is limited by the
+``max_page_sharing`` tunable. To increase the ratio ``max_page_sharing`` must
+be increased accordingly.
+
+--
+Izik Eidus,
+Hugh Dickins, 17 Nov 2009
diff --git a/Documentation/admin-guide/mm/numa_memory_policy.rst b/Documentation/admin-guide/mm/numa_memory_policy.rst
new file mode 100644
index 000000000..d78c5b315
--- /dev/null
+++ b/Documentation/admin-guide/mm/numa_memory_policy.rst
@@ -0,0 +1,495 @@
+.. _numa_memory_policy:
+
+==================
+NUMA Memory Policy
+==================
+
+What is NUMA Memory Policy?
+============================
+
+In the Linux kernel, "memory policy" determines from which node the kernel will
+allocate memory in a NUMA system or in an emulated NUMA system.  Linux has
+supported platforms with Non-Uniform Memory Access architectures since 2.4.?.
+The current memory policy support was added to Linux 2.6 around May 2004.  This
+document attempts to describe the concepts and APIs of the 2.6 memory policy
+support.
+
+Memory policies should not be confused with cpusets
+(``Documentation/cgroup-v1/cpusets.txt``)
+which is an administrative mechanism for restricting the nodes from which
+memory may be allocated by a set of processes. Memory policies are a
+programming interface that a NUMA-aware application can take advantage of.  When
+both cpusets and policies are applied to a task, the restrictions of the cpuset
+takes priority.  See :ref:`Memory Policies and cpusets <mem_pol_and_cpusets>`
+below for more details.
+
+Memory Policy Concepts
+======================
+
+Scope of Memory Policies
+------------------------
+
+The Linux kernel supports _scopes_ of memory policy, described here from
+most general to most specific:
+
+System Default Policy
+	this policy is "hard coded" into the kernel.  It is the policy
+	that governs all page allocations that aren't controlled by
+	one of the more specific policy scopes discussed below.  When
+	the system is "up and running", the system default policy will
+	use "local allocation" described below.  However, during boot
+	up, the system default policy will be set to interleave
+	allocations across all nodes with "sufficient" memory, so as
+	not to overload the initial boot node with boot-time
+	allocations.
+
+Task/Process Policy
+	this is an optional, per-task policy.  When defined for a
+	specific task, this policy controls all page allocations made
+	by or on behalf of the task that aren't controlled by a more
+	specific scope. If a task does not define a task policy, then
+	all page allocations that would have been controlled by the
+	task policy "fall back" to the System Default Policy.
+
+	The task policy applies to the entire address space of a task. Thus,
+	it is inheritable, and indeed is inherited, across both fork()
+	[clone() w/o the CLONE_VM flag] and exec*().  This allows a parent task
+	to establish the task policy for a child task exec()'d from an
+	executable image that has no awareness of memory policy.  See the
+	:ref:`Memory Policy APIs <memory_policy_apis>` section,
+	below, for an overview of the system call
+	that a task may use to set/change its task/process policy.
+
+	In a multi-threaded task, task policies apply only to the thread
+	[Linux kernel task] that installs the policy and any threads
+	subsequently created by that thread.  Any sibling threads existing
+	at the time a new task policy is installed retain their current
+	policy.
+
+	A task policy applies only to pages allocated after the policy is
+	installed.  Any pages already faulted in by the task when the task
+	changes its task policy remain where they were allocated based on
+	the policy at the time they were allocated.
+
+.. _vma_policy:
+
+VMA Policy
+	A "VMA" or "Virtual Memory Area" refers to a range of a task's
+	virtual address space.  A task may define a specific policy for a range
+	of its virtual address space.   See the
+	:ref:`Memory Policy APIs <memory_policy_apis>` section,
+	below, for an overview of the mbind() system call used to set a VMA
+	policy.
+
+	A VMA policy will govern the allocation of pages that back
+	this region of the address space.  Any regions of the task's
+	address space that don't have an explicit VMA policy will fall
+	back to the task policy, which may itself fall back to the
+	System Default Policy.
+
+	VMA policies have a few complicating details:
+
+	* VMA policy applies ONLY to anonymous pages.  These include
+	  pages allocated for anonymous segments, such as the task
+	  stack and heap, and any regions of the address space
+	  mmap()ed with the MAP_ANONYMOUS flag.  If a VMA policy is
+	  applied to a file mapping, it will be ignored if the mapping
+	  used the MAP_SHARED flag.  If the file mapping used the
+	  MAP_PRIVATE flag, the VMA policy will only be applied when
+	  an anonymous page is allocated on an attempt to write to the
+	  mapping-- i.e., at Copy-On-Write.
+
+	* VMA policies are shared between all tasks that share a
+	  virtual address space--a.k.a. threads--independent of when
+	  the policy is installed; and they are inherited across
+	  fork().  However, because VMA policies refer to a specific
+	  region of a task's address space, and because the address
+	  space is discarded and recreated on exec*(), VMA policies
+	  are NOT inheritable across exec().  Thus, only NUMA-aware
+	  applications may use VMA policies.
+
+	* A task may install a new VMA policy on a sub-range of a
+	  previously mmap()ed region.  When this happens, Linux splits
+	  the existing virtual memory area into 2 or 3 VMAs, each with
+	  it's own policy.
+
+	* By default, VMA policy applies only to pages allocated after
+	  the policy is installed.  Any pages already faulted into the
+	  VMA range remain where they were allocated based on the
+	  policy at the time they were allocated.  However, since
+	  2.6.16, Linux supports page migration via the mbind() system
+	  call, so that page contents can be moved to match a newly
+	  installed policy.
+
+Shared Policy
+	Conceptually, shared policies apply to "memory objects" mapped
+	shared into one or more tasks' distinct address spaces.  An
+	application installs shared policies the same way as VMA
+	policies--using the mbind() system call specifying a range of
+	virtual addresses that map the shared object.  However, unlike
+	VMA policies, which can be considered to be an attribute of a
+	range of a task's address space, shared policies apply
+	directly to the shared object.  Thus, all tasks that attach to
+	the object share the policy, and all pages allocated for the
+	shared object, by any task, will obey the shared policy.
+
+	As of 2.6.22, only shared memory segments, created by shmget() or
+	mmap(MAP_ANONYMOUS|MAP_SHARED), support shared policy.  When shared
+	policy support was added to Linux, the associated data structures were
+	added to hugetlbfs shmem segments.  At the time, hugetlbfs did not
+	support allocation at fault time--a.k.a lazy allocation--so hugetlbfs
+	shmem segments were never "hooked up" to the shared policy support.
+	Although hugetlbfs segments now support lazy allocation, their support
+	for shared policy has not been completed.
+
+	As mentioned above in :ref:`VMA policies <vma_policy>` section,
+	allocations of page cache pages for regular files mmap()ed
+	with MAP_SHARED ignore any VMA policy installed on the virtual
+	address range backed by the shared file mapping.  Rather,
+	shared page cache pages, including pages backing private
+	mappings that have not yet been written by the task, follow
+	task policy, if any, else System Default Policy.
+
+	The shared policy infrastructure supports different policies on subset
+	ranges of the shared object.  However, Linux still splits the VMA of
+	the task that installs the policy for each range of distinct policy.
+	Thus, different tasks that attach to a shared memory segment can have
+	different VMA configurations mapping that one shared object.  This
+	can be seen by examining the /proc/<pid>/numa_maps of tasks sharing
+	a shared memory region, when one task has installed shared policy on
+	one or more ranges of the region.
+
+Components of Memory Policies
+-----------------------------
+
+A NUMA memory policy consists of a "mode", optional mode flags, and
+an optional set of nodes.  The mode determines the behavior of the
+policy, the optional mode flags determine the behavior of the mode,
+and the optional set of nodes can be viewed as the arguments to the
+policy behavior.
+
+Internally, memory policies are implemented by a reference counted
+structure, struct mempolicy.  Details of this structure will be
+discussed in context, below, as required to explain the behavior.
+
+NUMA memory policy supports the following 4 behavioral modes:
+
+Default Mode--MPOL_DEFAULT
+	This mode is only used in the memory policy APIs.  Internally,
+	MPOL_DEFAULT is converted to the NULL memory policy in all
+	policy scopes.  Any existing non-default policy will simply be
+	removed when MPOL_DEFAULT is specified.  As a result,
+	MPOL_DEFAULT means "fall back to the next most specific policy
+	scope."
+
+	For example, a NULL or default task policy will fall back to the
+	system default policy.  A NULL or default vma policy will fall
+	back to the task policy.
+
+	When specified in one of the memory policy APIs, the Default mode
+	does not use the optional set of nodes.
+
+	It is an error for the set of nodes specified for this policy to
+	be non-empty.
+
+MPOL_BIND
+	This mode specifies that memory must come from the set of
+	nodes specified by the policy.  Memory will be allocated from
+	the node in the set with sufficient free memory that is
+	closest to the node where the allocation takes place.
+
+MPOL_PREFERRED
+	This mode specifies that the allocation should be attempted
+	from the single node specified in the policy.  If that
+	allocation fails, the kernel will search other nodes, in order
+	of increasing distance from the preferred node based on
+	information provided by the platform firmware.
+
+	Internally, the Preferred policy uses a single node--the
+	preferred_node member of struct mempolicy.  When the internal
+	mode flag MPOL_F_LOCAL is set, the preferred_node is ignored
+	and the policy is interpreted as local allocation.  "Local"
+	allocation policy can be viewed as a Preferred policy that
+	starts at the node containing the cpu where the allocation
+	takes place.
+
+	It is possible for the user to specify that local allocation
+	is always preferred by passing an empty nodemask with this
+	mode.  If an empty nodemask is passed, the policy cannot use
+	the MPOL_F_STATIC_NODES or MPOL_F_RELATIVE_NODES flags
+	described below.
+
+MPOL_INTERLEAVED
+	This mode specifies that page allocations be interleaved, on a
+	page granularity, across the nodes specified in the policy.
+	This mode also behaves slightly differently, based on the
+	context where it is used:
+
+	For allocation of anonymous pages and shared memory pages,
+	Interleave mode indexes the set of nodes specified by the
+	policy using the page offset of the faulting address into the
+	segment [VMA] containing the address modulo the number of
+	nodes specified by the policy.  It then attempts to allocate a
+	page, starting at the selected node, as if the node had been
+	specified by a Preferred policy or had been selected by a
+	local allocation.  That is, allocation will follow the per
+	node zonelist.
+
+	For allocation of page cache pages, Interleave mode indexes
+	the set of nodes specified by the policy using a node counter
+	maintained per task.  This counter wraps around to the lowest
+	specified node after it reaches the highest specified node.
+	This will tend to spread the pages out over the nodes
+	specified by the policy based on the order in which they are
+	allocated, rather than based on any page offset into an
+	address range or file.  During system boot up, the temporary
+	interleaved system default policy works in this mode.
+
+NUMA memory policy supports the following optional mode flags:
+
+MPOL_F_STATIC_NODES
+	This flag specifies that the nodemask passed by
+	the user should not be remapped if the task or VMA's set of allowed
+	nodes changes after the memory policy has been defined.
+
+	Without this flag, any time a mempolicy is rebound because of a
+	change in the set of allowed nodes, the node (Preferred) or
+	nodemask (Bind, Interleave) is remapped to the new set of
+	allowed nodes.  This may result in nodes being used that were
+	previously undesired.
+
+	With this flag, if the user-specified nodes overlap with the
+	nodes allowed by the task's cpuset, then the memory policy is
+	applied to their intersection.  If the two sets of nodes do not
+	overlap, the Default policy is used.
+
+	For example, consider a task that is attached to a cpuset with
+	mems 1-3 that sets an Interleave policy over the same set.  If
+	the cpuset's mems change to 3-5, the Interleave will now occur
+	over nodes 3, 4, and 5.  With this flag, however, since only node
+	3 is allowed from the user's nodemask, the "interleave" only
+	occurs over that node.  If no nodes from the user's nodemask are
+	now allowed, the Default behavior is used.
+
+	MPOL_F_STATIC_NODES cannot be combined with the
+	MPOL_F_RELATIVE_NODES flag.  It also cannot be used for
+	MPOL_PREFERRED policies that were created with an empty nodemask
+	(local allocation).
+
+MPOL_F_RELATIVE_NODES
+	This flag specifies that the nodemask passed
+	by the user will be mapped relative to the set of the task or VMA's
+	set of allowed nodes.  The kernel stores the user-passed nodemask,
+	and if the allowed nodes changes, then that original nodemask will
+	be remapped relative to the new set of allowed nodes.
+
+	Without this flag (and without MPOL_F_STATIC_NODES), anytime a
+	mempolicy is rebound because of a change in the set of allowed
+	nodes, the node (Preferred) or nodemask (Bind, Interleave) is
+	remapped to the new set of allowed nodes.  That remap may not
+	preserve the relative nature of the user's passed nodemask to its
+	set of allowed nodes upon successive rebinds: a nodemask of
+	1,3,5 may be remapped to 7-9 and then to 1-3 if the set of
+	allowed nodes is restored to its original state.
+
+	With this flag, the remap is done so that the node numbers from
+	the user's passed nodemask are relative to the set of allowed
+	nodes.  In other words, if nodes 0, 2, and 4 are set in the user's
+	nodemask, the policy will be effected over the first (and in the
+	Bind or Interleave case, the third and fifth) nodes in the set of
+	allowed nodes.  The nodemask passed by the user represents nodes
+	relative to task or VMA's set of allowed nodes.
+
+	If the user's nodemask includes nodes that are outside the range
+	of the new set of allowed nodes (for example, node 5 is set in
+	the user's nodemask when the set of allowed nodes is only 0-3),
+	then the remap wraps around to the beginning of the nodemask and,
+	if not already set, sets the node in the mempolicy nodemask.
+
+	For example, consider a task that is attached to a cpuset with
+	mems 2-5 that sets an Interleave policy over the same set with
+	MPOL_F_RELATIVE_NODES.  If the cpuset's mems change to 3-7, the
+	interleave now occurs over nodes 3,5-7.  If the cpuset's mems
+	then change to 0,2-3,5, then the interleave occurs over nodes
+	0,2-3,5.
+
+	Thanks to the consistent remapping, applications preparing
+	nodemasks to specify memory policies using this flag should
+	disregard their current, actual cpuset imposed memory placement
+	and prepare the nodemask as if they were always located on
+	memory nodes 0 to N-1, where N is the number of memory nodes the
+	policy is intended to manage.  Let the kernel then remap to the
+	set of memory nodes allowed by the task's cpuset, as that may
+	change over time.
+
+	MPOL_F_RELATIVE_NODES cannot be combined with the
+	MPOL_F_STATIC_NODES flag.  It also cannot be used for
+	MPOL_PREFERRED policies that were created with an empty nodemask
+	(local allocation).
+
+Memory Policy Reference Counting
+================================
+
+To resolve use/free races, struct mempolicy contains an atomic reference
+count field.  Internal interfaces, mpol_get()/mpol_put() increment and
+decrement this reference count, respectively.  mpol_put() will only free
+the structure back to the mempolicy kmem cache when the reference count
+goes to zero.
+
+When a new memory policy is allocated, its reference count is initialized
+to '1', representing the reference held by the task that is installing the
+new policy.  When a pointer to a memory policy structure is stored in another
+structure, another reference is added, as the task's reference will be dropped
+on completion of the policy installation.
+
+During run-time "usage" of the policy, we attempt to minimize atomic operations
+on the reference count, as this can lead to cache lines bouncing between cpus
+and NUMA nodes.  "Usage" here means one of the following:
+
+1) querying of the policy, either by the task itself [using the get_mempolicy()
+   API discussed below] or by another task using the /proc/<pid>/numa_maps
+   interface.
+
+2) examination of the policy to determine the policy mode and associated node
+   or node lists, if any, for page allocation.  This is considered a "hot
+   path".  Note that for MPOL_BIND, the "usage" extends across the entire
+   allocation process, which may sleep during page reclaimation, because the
+   BIND policy nodemask is used, by reference, to filter ineligible nodes.
+
+We can avoid taking an extra reference during the usages listed above as
+follows:
+
+1) we never need to get/free the system default policy as this is never
+   changed nor freed, once the system is up and running.
+
+2) for querying the policy, we do not need to take an extra reference on the
+   target task's task policy nor vma policies because we always acquire the
+   task's mm's mmap_sem for read during the query.  The set_mempolicy() and
+   mbind() APIs [see below] always acquire the mmap_sem for write when
+   installing or replacing task or vma policies.  Thus, there is no possibility
+   of a task or thread freeing a policy while another task or thread is
+   querying it.
+
+3) Page allocation usage of task or vma policy occurs in the fault path where
+   we hold them mmap_sem for read.  Again, because replacing the task or vma
+   policy requires that the mmap_sem be held for write, the policy can't be
+   freed out from under us while we're using it for page allocation.
+
+4) Shared policies require special consideration.  One task can replace a
+   shared memory policy while another task, with a distinct mmap_sem, is
+   querying or allocating a page based on the policy.  To resolve this
+   potential race, the shared policy infrastructure adds an extra reference
+   to the shared policy during lookup while holding a spin lock on the shared
+   policy management structure.  This requires that we drop this extra
+   reference when we're finished "using" the policy.  We must drop the
+   extra reference on shared policies in the same query/allocation paths
+   used for non-shared policies.  For this reason, shared policies are marked
+   as such, and the extra reference is dropped "conditionally"--i.e., only
+   for shared policies.
+
+   Because of this extra reference counting, and because we must lookup
+   shared policies in a tree structure under spinlock, shared policies are
+   more expensive to use in the page allocation path.  This is especially
+   true for shared policies on shared memory regions shared by tasks running
+   on different NUMA nodes.  This extra overhead can be avoided by always
+   falling back to task or system default policy for shared memory regions,
+   or by prefaulting the entire shared memory region into memory and locking
+   it down.  However, this might not be appropriate for all applications.
+
+.. _memory_policy_apis:
+
+Memory Policy APIs
+==================
+
+Linux supports 3 system calls for controlling memory policy.  These APIS
+always affect only the calling task, the calling task's address space, or
+some shared object mapped into the calling task's address space.
+
+.. note::
+   the headers that define these APIs and the parameter data types for
+   user space applications reside in a package that is not part of the
+   Linux kernel.  The kernel system call interfaces, with the 'sys\_'
+   prefix, are defined in <linux/syscalls.h>; the mode and flag
+   definitions are defined in <linux/mempolicy.h>.
+
+Set [Task] Memory Policy::
+
+	long set_mempolicy(int mode, const unsigned long *nmask,
+					unsigned long maxnode);
+
+Set's the calling task's "task/process memory policy" to mode
+specified by the 'mode' argument and the set of nodes defined by
+'nmask'.  'nmask' points to a bit mask of node ids containing at least
+'maxnode' ids.  Optional mode flags may be passed by combining the
+'mode' argument with the flag (for example: MPOL_INTERLEAVE |
+MPOL_F_STATIC_NODES).
+
+See the set_mempolicy(2) man page for more details
+
+
+Get [Task] Memory Policy or Related Information::
+
+	long get_mempolicy(int *mode,
+			   const unsigned long *nmask, unsigned long maxnode,
+			   void *addr, int flags);
+
+Queries the "task/process memory policy" of the calling task, or the
+policy or location of a specified virtual address, depending on the
+'flags' argument.
+
+See the get_mempolicy(2) man page for more details
+
+
+Install VMA/Shared Policy for a Range of Task's Address Space::
+
+	long mbind(void *start, unsigned long len, int mode,
+		   const unsigned long *nmask, unsigned long maxnode,
+		   unsigned flags);
+
+mbind() installs the policy specified by (mode, nmask, maxnodes) as a
+VMA policy for the range of the calling task's address space specified
+by the 'start' and 'len' arguments.  Additional actions may be
+requested via the 'flags' argument.
+
+See the mbind(2) man page for more details.
+
+Memory Policy Command Line Interface
+====================================
+
+Although not strictly part of the Linux implementation of memory policy,
+a command line tool, numactl(8), exists that allows one to:
+
++ set the task policy for a specified program via set_mempolicy(2), fork(2) and
+  exec(2)
+
++ set the shared policy for a shared memory segment via mbind(2)
+
+The numactl(8) tool is packaged with the run-time version of the library
+containing the memory policy system call wrappers.  Some distributions
+package the headers and compile-time libraries in a separate development
+package.
+
+.. _mem_pol_and_cpusets:
+
+Memory Policies and cpusets
+===========================
+
+Memory policies work within cpusets as described above.  For memory policies
+that require a node or set of nodes, the nodes are restricted to the set of
+nodes whose memories are allowed by the cpuset constraints.  If the nodemask
+specified for the policy contains nodes that are not allowed by the cpuset and
+MPOL_F_RELATIVE_NODES is not used, the intersection of the set of nodes
+specified for the policy and the set of nodes with memory is used.  If the
+result is the empty set, the policy is considered invalid and cannot be
+installed.  If MPOL_F_RELATIVE_NODES is used, the policy's nodes are mapped
+onto and folded into the task's set of allowed nodes as previously described.
+
+The interaction of memory policies and cpusets can be problematic when tasks
+in two cpusets share access to a memory region, such as shared memory segments
+created by shmget() of mmap() with the MAP_ANONYMOUS and MAP_SHARED flags, and
+any of the tasks install shared policy on the region, only nodes whose
+memories are allowed in both cpusets may be used in the policies.  Obtaining
+this information requires "stepping outside" the memory policy APIs to use the
+cpuset information and requires that one know in what cpusets other task might
+be attaching to the shared region.  Furthermore, if the cpusets' allowed
+memory sets are disjoint, "local" allocation is the only valid policy.
diff --git a/Documentation/admin-guide/mm/pagemap.rst b/Documentation/admin-guide/mm/pagemap.rst
new file mode 100644
index 000000000..3f7bade2c
--- /dev/null
+++ b/Documentation/admin-guide/mm/pagemap.rst
@@ -0,0 +1,204 @@
+.. _pagemap:
+
+=============================
+Examining Process Page Tables
+=============================
+
+pagemap is a new (as of 2.6.25) set of interfaces in the kernel that allow
+userspace programs to examine the page tables and related information by
+reading files in ``/proc``.
+
+There are four components to pagemap:
+
+ * ``/proc/pid/pagemap``.  This file lets a userspace process find out which
+   physical frame each virtual page is mapped to.  It contains one 64-bit
+   value for each virtual page, containing the following data (from
+   ``fs/proc/task_mmu.c``, above pagemap_read):
+
+    * Bits 0-54  page frame number (PFN) if present
+    * Bits 0-4   swap type if swapped
+    * Bits 5-54  swap offset if swapped
+    * Bit  55    pte is soft-dirty (see
+      :ref:`Documentation/admin-guide/mm/soft-dirty.rst <soft_dirty>`)
+    * Bit  56    page exclusively mapped (since 4.2)
+    * Bits 57-60 zero
+    * Bit  61    page is file-page or shared-anon (since 3.5)
+    * Bit  62    page swapped
+    * Bit  63    page present
+
+   Since Linux 4.0 only users with the CAP_SYS_ADMIN capability can get PFNs.
+   In 4.0 and 4.1 opens by unprivileged fail with -EPERM.  Starting from
+   4.2 the PFN field is zeroed if the user does not have CAP_SYS_ADMIN.
+   Reason: information about PFNs helps in exploiting Rowhammer vulnerability.
+
+   If the page is not present but in swap, then the PFN contains an
+   encoding of the swap file number and the page's offset into the
+   swap. Unmapped pages return a null PFN. This allows determining
+   precisely which pages are mapped (or in swap) and comparing mapped
+   pages between processes.
+
+   Efficient users of this interface will use ``/proc/pid/maps`` to
+   determine which areas of memory are actually mapped and llseek to
+   skip over unmapped regions.
+
+ * ``/proc/kpagecount``.  This file contains a 64-bit count of the number of
+   times each page is mapped, indexed by PFN.
+
+The page-types tool in the tools/vm directory can be used to query the
+number of times a page is mapped.
+
+ * ``/proc/kpageflags``.  This file contains a 64-bit set of flags for each
+   page, indexed by PFN.
+
+   The flags are (from ``fs/proc/page.c``, above kpageflags_read):
+
+    0. LOCKED
+    1. ERROR
+    2. REFERENCED
+    3. UPTODATE
+    4. DIRTY
+    5. LRU
+    6. ACTIVE
+    7. SLAB
+    8. WRITEBACK
+    9. RECLAIM
+    10. BUDDY
+    11. MMAP
+    12. ANON
+    13. SWAPCACHE
+    14. SWAPBACKED
+    15. COMPOUND_HEAD
+    16. COMPOUND_TAIL
+    17. HUGE
+    18. UNEVICTABLE
+    19. HWPOISON
+    20. NOPAGE
+    21. KSM
+    22. THP
+    23. BALLOON
+    24. ZERO_PAGE
+    25. IDLE
+
+ * ``/proc/kpagecgroup``.  This file contains a 64-bit inode number of the
+   memory cgroup each page is charged to, indexed by PFN. Only available when
+   CONFIG_MEMCG is set.
+
+Short descriptions to the page flags
+====================================
+
+0 - LOCKED
+   page is being locked for exclusive access, e.g. by undergoing read/write IO
+7 - SLAB
+   page is managed by the SLAB/SLOB/SLUB/SLQB kernel memory allocator
+   When compound page is used, SLUB/SLQB will only set this flag on the head
+   page; SLOB will not flag it at all.
+10 - BUDDY
+    a free memory block managed by the buddy system allocator
+    The buddy system organizes free memory in blocks of various orders.
+    An order N block has 2^N physically contiguous pages, with the BUDDY flag
+    set for and _only_ for the first page.
+15 - COMPOUND_HEAD
+    A compound page with order N consists of 2^N physically contiguous pages.
+    A compound page with order 2 takes the form of "HTTT", where H donates its
+    head page and T donates its tail page(s).  The major consumers of compound
+    pages are hugeTLB pages
+    (:ref:`Documentation/admin-guide/mm/hugetlbpage.rst <hugetlbpage>`),
+    the SLUB etc.  memory allocators and various device drivers.
+    However in this interface, only huge/giga pages are made visible
+    to end users.
+16 - COMPOUND_TAIL
+    A compound page tail (see description above).
+17 - HUGE
+    this is an integral part of a HugeTLB page
+19 - HWPOISON
+    hardware detected memory corruption on this page: don't touch the data!
+20 - NOPAGE
+    no page frame exists at the requested address
+21 - KSM
+    identical memory pages dynamically shared between one or more processes
+22 - THP
+    contiguous pages which construct transparent hugepages
+23 - BALLOON
+    balloon compaction page
+24 - ZERO_PAGE
+    zero page for pfn_zero or huge_zero page
+25 - IDLE
+    page has not been accessed since it was marked idle (see
+    :ref:`Documentation/admin-guide/mm/idle_page_tracking.rst <idle_page_tracking>`).
+    Note that this flag may be stale in case the page was accessed via
+    a PTE. To make sure the flag is up-to-date one has to read
+    ``/sys/kernel/mm/page_idle/bitmap`` first.
+
+IO related page flags
+---------------------
+
+1 - ERROR
+   IO error occurred
+3 - UPTODATE
+   page has up-to-date data
+   ie. for file backed page: (in-memory data revision >= on-disk one)
+4 - DIRTY
+   page has been written to, hence contains new data
+   i.e. for file backed page: (in-memory data revision >  on-disk one)
+8 - WRITEBACK
+   page is being synced to disk
+
+LRU related page flags
+----------------------
+
+5 - LRU
+   page is in one of the LRU lists
+6 - ACTIVE
+   page is in the active LRU list
+18 - UNEVICTABLE
+   page is in the unevictable (non-)LRU list It is somehow pinned and
+   not a candidate for LRU page reclaims, e.g. ramfs pages,
+   shmctl(SHM_LOCK) and mlock() memory segments
+2 - REFERENCED
+   page has been referenced since last LRU list enqueue/requeue
+9 - RECLAIM
+   page will be reclaimed soon after its pageout IO completed
+11 - MMAP
+   a memory mapped page
+12 - ANON
+   a memory mapped page that is not part of a file
+13 - SWAPCACHE
+   page is mapped to swap space, i.e. has an associated swap entry
+14 - SWAPBACKED
+   page is backed by swap/RAM
+
+The page-types tool in the tools/vm directory can be used to query the
+above flags.
+
+Using pagemap to do something useful
+====================================
+
+The general procedure for using pagemap to find out about a process' memory
+usage goes like this:
+
+ 1. Read ``/proc/pid/maps`` to determine which parts of the memory space are
+    mapped to what.
+ 2. Select the maps you are interested in -- all of them, or a particular
+    library, or the stack or the heap, etc.
+ 3. Open ``/proc/pid/pagemap`` and seek to the pages you would like to examine.
+ 4. Read a u64 for each page from pagemap.
+ 5. Open ``/proc/kpagecount`` and/or ``/proc/kpageflags``.  For each PFN you
+    just read, seek to that entry in the file, and read the data you want.
+
+For example, to find the "unique set size" (USS), which is the amount of
+memory that a process is using that is not shared with any other process,
+you can go through every map in the process, find the PFNs, look those up
+in kpagecount, and tally up the number of pages that are only referenced
+once.
+
+Other notes
+===========
+
+Reading from any of the files will return -EINVAL if you are not starting
+the read on an 8-byte boundary (e.g., if you sought an odd number of bytes
+into the file), or if the size of the read is not a multiple of 8 bytes.
+
+Before Linux 3.11 pagemap bits 55-60 were used for "page-shift" (which is
+always 12 at most architectures). Since Linux 3.11 their meaning changes
+after first clear of soft-dirty bits. Since Linux 4.2 they are used for
+flags unconditionally.
diff --git a/Documentation/admin-guide/mm/soft-dirty.rst b/Documentation/admin-guide/mm/soft-dirty.rst
new file mode 100644
index 000000000..cb0cfd667
--- /dev/null
+++ b/Documentation/admin-guide/mm/soft-dirty.rst
@@ -0,0 +1,47 @@
+.. _soft_dirty:
+
+===============
+Soft-Dirty PTEs
+===============
+
+The soft-dirty is a bit on a PTE which helps to track which pages a task
+writes to. In order to do this tracking one should
+
+  1. Clear soft-dirty bits from the task's PTEs.
+
+     This is done by writing "4" into the ``/proc/PID/clear_refs`` file of the
+     task in question.
+
+  2. Wait some time.
+
+  3. Read soft-dirty bits from the PTEs.
+
+     This is done by reading from the ``/proc/PID/pagemap``. The bit 55 of the
+     64-bit qword is the soft-dirty one. If set, the respective PTE was
+     written to since step 1.
+
+
+Internally, to do this tracking, the writable bit is cleared from PTEs
+when the soft-dirty bit is cleared. So, after this, when the task tries to
+modify a page at some virtual address the #PF occurs and the kernel sets
+the soft-dirty bit on the respective PTE.
+
+Note, that although all the task's address space is marked as r/o after the
+soft-dirty bits clear, the #PF-s that occur after that are processed fast.
+This is so, since the pages are still mapped to physical memory, and thus all
+the kernel does is finds this fact out and puts both writable and soft-dirty
+bits on the PTE.
+
+While in most cases tracking memory changes by #PF-s is more than enough
+there is still a scenario when we can lose soft dirty bits -- a task
+unmaps a previously mapped memory region and then maps a new one at exactly
+the same place. When unmap is called, the kernel internally clears PTE values
+including soft dirty bits. To notify user space application about such
+memory region renewal the kernel always marks new memory regions (and
+expanded regions) as soft dirty.
+
+This feature is actively used by the checkpoint-restore project. You
+can find more details about it on http://criu.org
+
+
+-- Pavel Emelyanov, Apr 9, 2013
diff --git a/Documentation/admin-guide/mm/transhuge.rst b/Documentation/admin-guide/mm/transhuge.rst
new file mode 100644
index 000000000..7ab93a840
--- /dev/null
+++ b/Documentation/admin-guide/mm/transhuge.rst
@@ -0,0 +1,418 @@
+.. _admin_guide_transhuge:
+
+============================
+Transparent Hugepage Support
+============================
+
+Objective
+=========
+
+Performance critical computing applications dealing with large memory
+working sets are already running on top of libhugetlbfs and in turn
+hugetlbfs. Transparent HugePage Support (THP) is an alternative mean of
+using huge pages for the backing of virtual memory with huge pages
+that supports the automatic promotion and demotion of page sizes and
+without the shortcomings of hugetlbfs.
+
+Currently THP only works for anonymous memory mappings and tmpfs/shmem.
+But in the future it can expand to other filesystems.
+
+.. note::
+   in the examples below we presume that the basic page size is 4K and
+   the huge page size is 2M, although the actual numbers may vary
+   depending on the CPU architecture.
+
+The reason applications are running faster is because of two
+factors. The first factor is almost completely irrelevant and it's not
+of significant interest because it'll also have the downside of
+requiring larger clear-page copy-page in page faults which is a
+potentially negative effect. The first factor consists in taking a
+single page fault for each 2M virtual region touched by userland (so
+reducing the enter/exit kernel frequency by a 512 times factor). This
+only matters the first time the memory is accessed for the lifetime of
+a memory mapping. The second long lasting and much more important
+factor will affect all subsequent accesses to the memory for the whole
+runtime of the application. The second factor consist of two
+components:
+
+1) the TLB miss will run faster (especially with virtualization using
+   nested pagetables but almost always also on bare metal without
+   virtualization)
+
+2) a single TLB entry will be mapping a much larger amount of virtual
+   memory in turn reducing the number of TLB misses. With
+   virtualization and nested pagetables the TLB can be mapped of
+   larger size only if both KVM and the Linux guest are using
+   hugepages but a significant speedup already happens if only one of
+   the two is using hugepages just because of the fact the TLB miss is
+   going to run faster.
+
+THP can be enabled system wide or restricted to certain tasks or even
+memory ranges inside task's address space. Unless THP is completely
+disabled, there is ``khugepaged`` daemon that scans memory and
+collapses sequences of basic pages into huge pages.
+
+The THP behaviour is controlled via :ref:`sysfs <thp_sysfs>`
+interface and using madivse(2) and prctl(2) system calls.
+
+Transparent Hugepage Support maximizes the usefulness of free memory
+if compared to the reservation approach of hugetlbfs by allowing all
+unused memory to be used as cache or other movable (or even unmovable
+entities). It doesn't require reservation to prevent hugepage
+allocation failures to be noticeable from userland. It allows paging
+and all other advanced VM features to be available on the
+hugepages. It requires no modifications for applications to take
+advantage of it.
+
+Applications however can be further optimized to take advantage of
+this feature, like for example they've been optimized before to avoid
+a flood of mmap system calls for every malloc(4k). Optimizing userland
+is by far not mandatory and khugepaged already can take care of long
+lived page allocations even for hugepage unaware applications that
+deals with large amounts of memory.
+
+In certain cases when hugepages are enabled system wide, application
+may end up allocating more memory resources. An application may mmap a
+large region but only touch 1 byte of it, in that case a 2M page might
+be allocated instead of a 4k page for no good. This is why it's
+possible to disable hugepages system-wide and to only have them inside
+MADV_HUGEPAGE madvise regions.
+
+Embedded systems should enable hugepages only inside madvise regions
+to eliminate any risk of wasting any precious byte of memory and to
+only run faster.
+
+Applications that gets a lot of benefit from hugepages and that don't
+risk to lose memory by using hugepages, should use
+madvise(MADV_HUGEPAGE) on their critical mmapped regions.
+
+.. _thp_sysfs:
+
+sysfs
+=====
+
+Global THP controls
+-------------------
+
+Transparent Hugepage Support for anonymous memory can be entirely disabled
+(mostly for debugging purposes) or only enabled inside MADV_HUGEPAGE
+regions (to avoid the risk of consuming more memory resources) or enabled
+system wide. This can be achieved with one of::
+
+	echo always >/sys/kernel/mm/transparent_hugepage/enabled
+	echo madvise >/sys/kernel/mm/transparent_hugepage/enabled
+	echo never >/sys/kernel/mm/transparent_hugepage/enabled
+
+It's also possible to limit defrag efforts in the VM to generate
+anonymous hugepages in case they're not immediately free to madvise
+regions or to never try to defrag memory and simply fallback to regular
+pages unless hugepages are immediately available. Clearly if we spend CPU
+time to defrag memory, we would expect to gain even more by the fact we
+use hugepages later instead of regular pages. This isn't always
+guaranteed, but it may be more likely in case the allocation is for a
+MADV_HUGEPAGE region.
+
+::
+
+	echo always >/sys/kernel/mm/transparent_hugepage/defrag
+	echo defer >/sys/kernel/mm/transparent_hugepage/defrag
+	echo defer+madvise >/sys/kernel/mm/transparent_hugepage/defrag
+	echo madvise >/sys/kernel/mm/transparent_hugepage/defrag
+	echo never >/sys/kernel/mm/transparent_hugepage/defrag
+
+always
+	means that an application requesting THP will stall on
+	allocation failure and directly reclaim pages and compact
+	memory in an effort to allocate a THP immediately. This may be
+	desirable for virtual machines that benefit heavily from THP
+	use and are willing to delay the VM start to utilise them.
+
+defer
+	means that an application will wake kswapd in the background
+	to reclaim pages and wake kcompactd to compact memory so that
+	THP is available in the near future. It's the responsibility
+	of khugepaged to then install the THP pages later.
+
+defer+madvise
+	will enter direct reclaim and compaction like ``always``, but
+	only for regions that have used madvise(MADV_HUGEPAGE); all
+	other regions will wake kswapd in the background to reclaim
+	pages and wake kcompactd to compact memory so that THP is
+	available in the near future.
+
+madvise
+	will enter direct reclaim like ``always`` but only for regions
+	that are have used madvise(MADV_HUGEPAGE). This is the default
+	behaviour.
+
+never
+	should be self-explanatory.
+
+By default kernel tries to use huge zero page on read page fault to
+anonymous mapping. It's possible to disable huge zero page by writing 0
+or enable it back by writing 1::
+
+	echo 0 >/sys/kernel/mm/transparent_hugepage/use_zero_page
+	echo 1 >/sys/kernel/mm/transparent_hugepage/use_zero_page
+
+Some userspace (such as a test program, or an optimized memory allocation
+library) may want to know the size (in bytes) of a transparent hugepage::
+
+	cat /sys/kernel/mm/transparent_hugepage/hpage_pmd_size
+
+khugepaged will be automatically started when
+transparent_hugepage/enabled is set to "always" or "madvise, and it'll
+be automatically shutdown if it's set to "never".
+
+Khugepaged controls
+-------------------
+
+khugepaged runs usually at low frequency so while one may not want to
+invoke defrag algorithms synchronously during the page faults, it
+should be worth invoking defrag at least in khugepaged. However it's
+also possible to disable defrag in khugepaged by writing 0 or enable
+defrag in khugepaged by writing 1::
+
+	echo 0 >/sys/kernel/mm/transparent_hugepage/khugepaged/defrag
+	echo 1 >/sys/kernel/mm/transparent_hugepage/khugepaged/defrag
+
+You can also control how many pages khugepaged should scan at each
+pass::
+
+	/sys/kernel/mm/transparent_hugepage/khugepaged/pages_to_scan
+
+and how many milliseconds to wait in khugepaged between each pass (you
+can set this to 0 to run khugepaged at 100% utilization of one core)::
+
+	/sys/kernel/mm/transparent_hugepage/khugepaged/scan_sleep_millisecs
+
+and how many milliseconds to wait in khugepaged if there's an hugepage
+allocation failure to throttle the next allocation attempt::
+
+	/sys/kernel/mm/transparent_hugepage/khugepaged/alloc_sleep_millisecs
+
+The khugepaged progress can be seen in the number of pages collapsed::
+
+	/sys/kernel/mm/transparent_hugepage/khugepaged/pages_collapsed
+
+for each pass::
+
+	/sys/kernel/mm/transparent_hugepage/khugepaged/full_scans
+
+``max_ptes_none`` specifies how many extra small pages (that are
+not already mapped) can be allocated when collapsing a group
+of small pages into one large page::
+
+	/sys/kernel/mm/transparent_hugepage/khugepaged/max_ptes_none
+
+A higher value leads to use additional memory for programs.
+A lower value leads to gain less thp performance. Value of
+max_ptes_none can waste cpu time very little, you can
+ignore it.
+
+``max_ptes_swap`` specifies how many pages can be brought in from
+swap when collapsing a group of pages into a transparent huge page::
+
+	/sys/kernel/mm/transparent_hugepage/khugepaged/max_ptes_swap
+
+A higher value can cause excessive swap IO and waste
+memory. A lower value can prevent THPs from being
+collapsed, resulting fewer pages being collapsed into
+THPs, and lower memory access performance.
+
+Boot parameter
+==============
+
+You can change the sysfs boot time defaults of Transparent Hugepage
+Support by passing the parameter ``transparent_hugepage=always`` or
+``transparent_hugepage=madvise`` or ``transparent_hugepage=never``
+to the kernel command line.
+
+Hugepages in tmpfs/shmem
+========================
+
+You can control hugepage allocation policy in tmpfs with mount option
+``huge=``. It can have following values:
+
+always
+    Attempt to allocate huge pages every time we need a new page;
+
+never
+    Do not allocate huge pages;
+
+within_size
+    Only allocate huge page if it will be fully within i_size.
+    Also respect fadvise()/madvise() hints;
+
+advise
+    Only allocate huge pages if requested with fadvise()/madvise();
+
+The default policy is ``never``.
+
+``mount -o remount,huge= /mountpoint`` works fine after mount: remounting
+``huge=never`` will not attempt to break up huge pages at all, just stop more
+from being allocated.
+
+There's also sysfs knob to control hugepage allocation policy for internal
+shmem mount: /sys/kernel/mm/transparent_hugepage/shmem_enabled. The mount
+is used for SysV SHM, memfds, shared anonymous mmaps (of /dev/zero or
+MAP_ANONYMOUS), GPU drivers' DRM objects, Ashmem.
+
+In addition to policies listed above, shmem_enabled allows two further
+values:
+
+deny
+    For use in emergencies, to force the huge option off from
+    all mounts;
+force
+    Force the huge option on for all - very useful for testing;
+
+Need of application restart
+===========================
+
+The transparent_hugepage/enabled values and tmpfs mount option only affect
+future behavior. So to make them effective you need to restart any
+application that could have been using hugepages. This also applies to the
+regions registered in khugepaged.
+
+Monitoring usage
+================
+
+The number of anonymous transparent huge pages currently used by the
+system is available by reading the AnonHugePages field in ``/proc/meminfo``.
+To identify what applications are using anonymous transparent huge pages,
+it is necessary to read ``/proc/PID/smaps`` and count the AnonHugePages fields
+for each mapping.
+
+The number of file transparent huge pages mapped to userspace is available
+by reading ShmemPmdMapped and ShmemHugePages fields in ``/proc/meminfo``.
+To identify what applications are mapping file transparent huge pages, it
+is necessary to read ``/proc/PID/smaps`` and count the FileHugeMapped fields
+for each mapping.
+
+Note that reading the smaps file is expensive and reading it
+frequently will incur overhead.
+
+There are a number of counters in ``/proc/vmstat`` that may be used to
+monitor how successfully the system is providing huge pages for use.
+
+thp_fault_alloc
+	is incremented every time a huge page is successfully
+	allocated to handle a page fault. This applies to both the
+	first time a page is faulted and for COW faults.
+
+thp_collapse_alloc
+	is incremented by khugepaged when it has found
+	a range of pages to collapse into one huge page and has
+	successfully allocated a new huge page to store the data.
+
+thp_fault_fallback
+	is incremented if a page fault fails to allocate
+	a huge page and instead falls back to using small pages.
+
+thp_collapse_alloc_failed
+	is incremented if khugepaged found a range
+	of pages that should be collapsed into one huge page but failed
+	the allocation.
+
+thp_file_alloc
+	is incremented every time a file huge page is successfully
+	allocated.
+
+thp_file_mapped
+	is incremented every time a file huge page is mapped into
+	user address space.
+
+thp_split_page
+	is incremented every time a huge page is split into base
+	pages. This can happen for a variety of reasons but a common
+	reason is that a huge page is old and is being reclaimed.
+	This action implies splitting all PMD the page mapped with.
+
+thp_split_page_failed
+	is incremented if kernel fails to split huge
+	page. This can happen if the page was pinned by somebody.
+
+thp_deferred_split_page
+	is incremented when a huge page is put onto split
+	queue. This happens when a huge page is partially unmapped and
+	splitting it would free up some memory. Pages on split queue are
+	going to be split under memory pressure.
+
+thp_split_pmd
+	is incremented every time a PMD split into table of PTEs.
+	This can happen, for instance, when application calls mprotect() or
+	munmap() on part of huge page. It doesn't split huge page, only
+	page table entry.
+
+thp_zero_page_alloc
+	is incremented every time a huge zero page is
+	successfully allocated. It includes allocations which where
+	dropped due race with other allocation. Note, it doesn't count
+	every map of the huge zero page, only its allocation.
+
+thp_zero_page_alloc_failed
+	is incremented if kernel fails to allocate
+	huge zero page and falls back to using small pages.
+
+thp_swpout
+	is incremented every time a huge page is swapout in one
+	piece without splitting.
+
+thp_swpout_fallback
+	is incremented if a huge page has to be split before swapout.
+	Usually because failed to allocate some continuous swap space
+	for the huge page.
+
+As the system ages, allocating huge pages may be expensive as the
+system uses memory compaction to copy data around memory to free a
+huge page for use. There are some counters in ``/proc/vmstat`` to help
+monitor this overhead.
+
+compact_stall
+	is incremented every time a process stalls to run
+	memory compaction so that a huge page is free for use.
+
+compact_success
+	is incremented if the system compacted memory and
+	freed a huge page for use.
+
+compact_fail
+	is incremented if the system tries to compact memory
+	but failed.
+
+compact_pages_moved
+	is incremented each time a page is moved. If
+	this value is increasing rapidly, it implies that the system
+	is copying a lot of data to satisfy the huge page allocation.
+	It is possible that the cost of copying exceeds any savings
+	from reduced TLB misses.
+
+compact_pagemigrate_failed
+	is incremented when the underlying mechanism
+	for moving a page failed.
+
+compact_blocks_moved
+	is incremented each time memory compaction examines
+	a huge page aligned range of pages.
+
+It is possible to establish how long the stalls were using the function
+tracer to record how long was spent in __alloc_pages_nodemask and
+using the mm_page_alloc tracepoint to identify which allocations were
+for huge pages.
+
+Optimizing the applications
+===========================
+
+To be guaranteed that the kernel will map a 2M page immediately in any
+memory region, the mmap region has to be hugepage naturally
+aligned. posix_memalign() can provide that guarantee.
+
+Hugetlbfs
+=========
+
+You can use hugetlbfs on a kernel that has transparent hugepage
+support enabled just fine as always. No difference can be noted in
+hugetlbfs other than there will be less overall fragmentation. All
+usual features belonging to hugetlbfs are preserved and
+unaffected. libhugetlbfs will also work fine as usual.
diff --git a/Documentation/admin-guide/mm/userfaultfd.rst b/Documentation/admin-guide/mm/userfaultfd.rst
new file mode 100644
index 000000000..5048cf661
--- /dev/null
+++ b/Documentation/admin-guide/mm/userfaultfd.rst
@@ -0,0 +1,241 @@
+.. _userfaultfd:
+
+===========
+Userfaultfd
+===========
+
+Objective
+=========
+
+Userfaults allow the implementation of on-demand paging from userland
+and more generally they allow userland to take control of various
+memory page faults, something otherwise only the kernel code could do.
+
+For example userfaults allows a proper and more optimal implementation
+of the PROT_NONE+SIGSEGV trick.
+
+Design
+======
+
+Userfaults are delivered and resolved through the userfaultfd syscall.
+
+The userfaultfd (aside from registering and unregistering virtual
+memory ranges) provides two primary functionalities:
+
+1) read/POLLIN protocol to notify a userland thread of the faults
+   happening
+
+2) various UFFDIO_* ioctls that can manage the virtual memory regions
+   registered in the userfaultfd that allows userland to efficiently
+   resolve the userfaults it receives via 1) or to manage the virtual
+   memory in the background
+
+The real advantage of userfaults if compared to regular virtual memory
+management of mremap/mprotect is that the userfaults in all their
+operations never involve heavyweight structures like vmas (in fact the
+userfaultfd runtime load never takes the mmap_sem for writing).
+
+Vmas are not suitable for page- (or hugepage) granular fault tracking
+when dealing with virtual address spaces that could span
+Terabytes. Too many vmas would be needed for that.
+
+The userfaultfd once opened by invoking the syscall, can also be
+passed using unix domain sockets to a manager process, so the same
+manager process could handle the userfaults of a multitude of
+different processes without them being aware about what is going on
+(well of course unless they later try to use the userfaultfd
+themselves on the same region the manager is already tracking, which
+is a corner case that would currently return -EBUSY).
+
+API
+===
+
+When first opened the userfaultfd must be enabled invoking the
+UFFDIO_API ioctl specifying a uffdio_api.api value set to UFFD_API (or
+a later API version) which will specify the read/POLLIN protocol
+userland intends to speak on the UFFD and the uffdio_api.features
+userland requires. The UFFDIO_API ioctl if successful (i.e. if the
+requested uffdio_api.api is spoken also by the running kernel and the
+requested features are going to be enabled) will return into
+uffdio_api.features and uffdio_api.ioctls two 64bit bitmasks of
+respectively all the available features of the read(2) protocol and
+the generic ioctl available.
+
+The uffdio_api.features bitmask returned by the UFFDIO_API ioctl
+defines what memory types are supported by the userfaultfd and what
+events, except page fault notifications, may be generated.
+
+If the kernel supports registering userfaultfd ranges on hugetlbfs
+virtual memory areas, UFFD_FEATURE_MISSING_HUGETLBFS will be set in
+uffdio_api.features. Similarly, UFFD_FEATURE_MISSING_SHMEM will be
+set if the kernel supports registering userfaultfd ranges on shared
+memory (covering all shmem APIs, i.e. tmpfs, IPCSHM, /dev/zero
+MAP_SHARED, memfd_create, etc).
+
+The userland application that wants to use userfaultfd with hugetlbfs
+or shared memory need to set the corresponding flag in
+uffdio_api.features to enable those features.
+
+If the userland desires to receive notifications for events other than
+page faults, it has to verify that uffdio_api.features has appropriate
+UFFD_FEATURE_EVENT_* bits set. These events are described in more
+detail below in "Non-cooperative userfaultfd" section.
+
+Once the userfaultfd has been enabled the UFFDIO_REGISTER ioctl should
+be invoked (if present in the returned uffdio_api.ioctls bitmask) to
+register a memory range in the userfaultfd by setting the
+uffdio_register structure accordingly. The uffdio_register.mode
+bitmask will specify to the kernel which kind of faults to track for
+the range (UFFDIO_REGISTER_MODE_MISSING would track missing
+pages). The UFFDIO_REGISTER ioctl will return the
+uffdio_register.ioctls bitmask of ioctls that are suitable to resolve
+userfaults on the range registered. Not all ioctls will necessarily be
+supported for all memory types depending on the underlying virtual
+memory backend (anonymous memory vs tmpfs vs real filebacked
+mappings).
+
+Userland can use the uffdio_register.ioctls to manage the virtual
+address space in the background (to add or potentially also remove
+memory from the userfaultfd registered range). This means a userfault
+could be triggering just before userland maps in the background the
+user-faulted page.
+
+The primary ioctl to resolve userfaults is UFFDIO_COPY. That
+atomically copies a page into the userfault registered range and wakes
+up the blocked userfaults (unless uffdio_copy.mode &
+UFFDIO_COPY_MODE_DONTWAKE is set). Other ioctl works similarly to
+UFFDIO_COPY. They're atomic as in guaranteeing that nothing can see an
+half copied page since it'll keep userfaulting until the copy has
+finished.
+
+QEMU/KVM
+========
+
+QEMU/KVM is using the userfaultfd syscall to implement postcopy live
+migration. Postcopy live migration is one form of memory
+externalization consisting of a virtual machine running with part or
+all of its memory residing on a different node in the cloud. The
+userfaultfd abstraction is generic enough that not a single line of
+KVM kernel code had to be modified in order to add postcopy live
+migration to QEMU.
+
+Guest async page faults, FOLL_NOWAIT and all other GUP features work
+just fine in combination with userfaults. Userfaults trigger async
+page faults in the guest scheduler so those guest processes that
+aren't waiting for userfaults (i.e. network bound) can keep running in
+the guest vcpus.
+
+It is generally beneficial to run one pass of precopy live migration
+just before starting postcopy live migration, in order to avoid
+generating userfaults for readonly guest regions.
+
+The implementation of postcopy live migration currently uses one
+single bidirectional socket but in the future two different sockets
+will be used (to reduce the latency of the userfaults to the minimum
+possible without having to decrease /proc/sys/net/ipv4/tcp_wmem).
+
+The QEMU in the source node writes all pages that it knows are missing
+in the destination node, into the socket, and the migration thread of
+the QEMU running in the destination node runs UFFDIO_COPY|ZEROPAGE
+ioctls on the userfaultfd in order to map the received pages into the
+guest (UFFDIO_ZEROCOPY is used if the source page was a zero page).
+
+A different postcopy thread in the destination node listens with
+poll() to the userfaultfd in parallel. When a POLLIN event is
+generated after a userfault triggers, the postcopy thread read() from
+the userfaultfd and receives the fault address (or -EAGAIN in case the
+userfault was already resolved and waken by a UFFDIO_COPY|ZEROPAGE run
+by the parallel QEMU migration thread).
+
+After the QEMU postcopy thread (running in the destination node) gets
+the userfault address it writes the information about the missing page
+into the socket. The QEMU source node receives the information and
+roughly "seeks" to that page address and continues sending all
+remaining missing pages from that new page offset. Soon after that
+(just the time to flush the tcp_wmem queue through the network) the
+migration thread in the QEMU running in the destination node will
+receive the page that triggered the userfault and it'll map it as
+usual with the UFFDIO_COPY|ZEROPAGE (without actually knowing if it
+was spontaneously sent by the source or if it was an urgent page
+requested through a userfault).
+
+By the time the userfaults start, the QEMU in the destination node
+doesn't need to keep any per-page state bitmap relative to the live
+migration around and a single per-page bitmap has to be maintained in
+the QEMU running in the source node to know which pages are still
+missing in the destination node. The bitmap in the source node is
+checked to find which missing pages to send in round robin and we seek
+over it when receiving incoming userfaults. After sending each page of
+course the bitmap is updated accordingly. It's also useful to avoid
+sending the same page twice (in case the userfault is read by the
+postcopy thread just before UFFDIO_COPY|ZEROPAGE runs in the migration
+thread).
+
+Non-cooperative userfaultfd
+===========================
+
+When the userfaultfd is monitored by an external manager, the manager
+must be able to track changes in the process virtual memory
+layout. Userfaultfd can notify the manager about such changes using
+the same read(2) protocol as for the page fault notifications. The
+manager has to explicitly enable these events by setting appropriate
+bits in uffdio_api.features passed to UFFDIO_API ioctl:
+
+UFFD_FEATURE_EVENT_FORK
+	enable userfaultfd hooks for fork(). When this feature is
+	enabled, the userfaultfd context of the parent process is
+	duplicated into the newly created process. The manager
+	receives UFFD_EVENT_FORK with file descriptor of the new
+	userfaultfd context in the uffd_msg.fork.
+
+UFFD_FEATURE_EVENT_REMAP
+	enable notifications about mremap() calls. When the
+	non-cooperative process moves a virtual memory area to a
+	different location, the manager will receive
+	UFFD_EVENT_REMAP. The uffd_msg.remap will contain the old and
+	new addresses of the area and its original length.
+
+UFFD_FEATURE_EVENT_REMOVE
+	enable notifications about madvise(MADV_REMOVE) and
+	madvise(MADV_DONTNEED) calls. The event UFFD_EVENT_REMOVE will
+	be generated upon these calls to madvise. The uffd_msg.remove
+	will contain start and end addresses of the removed area.
+
+UFFD_FEATURE_EVENT_UNMAP
+	enable notifications about memory unmapping. The manager will
+	get UFFD_EVENT_UNMAP with uffd_msg.remove containing start and
+	end addresses of the unmapped area.
+
+Although the UFFD_FEATURE_EVENT_REMOVE and UFFD_FEATURE_EVENT_UNMAP
+are pretty similar, they quite differ in the action expected from the
+userfaultfd manager. In the former case, the virtual memory is
+removed, but the area is not, the area remains monitored by the
+userfaultfd, and if a page fault occurs in that area it will be
+delivered to the manager. The proper resolution for such page fault is
+to zeromap the faulting address. However, in the latter case, when an
+area is unmapped, either explicitly (with munmap() system call), or
+implicitly (e.g. during mremap()), the area is removed and in turn the
+userfaultfd context for such area disappears too and the manager will
+not get further userland page faults from the removed area. Still, the
+notification is required in order to prevent manager from using
+UFFDIO_COPY on the unmapped area.
+
+Unlike userland page faults which have to be synchronous and require
+explicit or implicit wakeup, all the events are delivered
+asynchronously and the non-cooperative process resumes execution as
+soon as manager executes read(). The userfaultfd manager should
+carefully synchronize calls to UFFDIO_COPY with the events
+processing. To aid the synchronization, the UFFDIO_COPY ioctl will
+return -ENOSPC when the monitored process exits at the time of
+UFFDIO_COPY, and -ENOENT, when the non-cooperative process has changed
+its virtual memory layout simultaneously with outstanding UFFDIO_COPY
+operation.
+
+The current asynchronous model of the event delivery is optimal for
+single threaded non-cooperative userfaultfd manager implementations. A
+synchronous event delivery model can be added later as a new
+userfaultfd feature to facilitate multithreading enhancements of the
+non cooperative manager, for example to allow UFFDIO_COPY ioctls to
+run in parallel to the event reception. Single threaded
+implementations should continue to use the current async event
+delivery model instead.