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-rw-r--r--Documentation/vm/.gitignore3
-rw-r--r--Documentation/vm/active_mm.rst91
-rw-r--r--Documentation/vm/arch_pgtable_helpers.rst258
-rw-r--r--Documentation/vm/balance.rst102
-rw-r--r--Documentation/vm/cleancache.rst296
-rw-r--r--Documentation/vm/free_page_reporting.rst40
-rw-r--r--Documentation/vm/frontswap.rst293
-rw-r--r--Documentation/vm/highmem.rst147
-rw-r--r--Documentation/vm/hmm.rst435
-rw-r--r--Documentation/vm/hugetlbfs_reserv.rst596
-rw-r--r--Documentation/vm/hwpoison.rst186
-rw-r--r--Documentation/vm/index.rst55
-rw-r--r--Documentation/vm/ksm.rst87
-rw-r--r--Documentation/vm/memory-model.rst218
-rw-r--r--Documentation/vm/mmu_notifier.rst99
-rw-r--r--Documentation/vm/numa.rst150
-rw-r--r--Documentation/vm/overcommit-accounting.rst87
-rw-r--r--Documentation/vm/page_frags.rst45
-rw-r--r--Documentation/vm/page_migration.rst288
-rw-r--r--Documentation/vm/page_owner.rst89
-rw-r--r--Documentation/vm/remap_file_pages.rst33
-rw-r--r--Documentation/vm/slub.rst388
-rw-r--r--Documentation/vm/split_page_table_lock.rst100
-rw-r--r--Documentation/vm/swap_numa.rst80
-rw-r--r--Documentation/vm/transhuge.rst192
-rw-r--r--Documentation/vm/unevictable-lru.rst618
-rw-r--r--Documentation/vm/z3fold.rst30
-rw-r--r--Documentation/vm/zsmalloc.rst82
-rw-r--r--Documentation/vm/zswap.rst152
29 files changed, 5240 insertions, 0 deletions
diff --git a/Documentation/vm/.gitignore b/Documentation/vm/.gitignore
new file mode 100644
index 000000000..bc74f5643
--- /dev/null
+++ b/Documentation/vm/.gitignore
@@ -0,0 +1,3 @@
+# SPDX-License-Identifier: GPL-2.0-only
+page-types
+slabinfo
diff --git a/Documentation/vm/active_mm.rst b/Documentation/vm/active_mm.rst
new file mode 100644
index 000000000..6f8269c28
--- /dev/null
+++ b/Documentation/vm/active_mm.rst
@@ -0,0 +1,91 @@
+.. _active_mm:
+
+=========
+Active MM
+=========
+
+::
+
+ List: linux-kernel
+ Subject: Re: active_mm
+ From: Linus Torvalds <torvalds () transmeta ! com>
+ Date: 1999-07-30 21:36:24
+
+ Cc'd to linux-kernel, because I don't write explanations all that often,
+ and when I do I feel better about more people reading them.
+
+ On Fri, 30 Jul 1999, David Mosberger wrote:
+ >
+ > Is there a brief description someplace on how "mm" vs. "active_mm" in
+ > the task_struct are supposed to be used? (My apologies if this was
+ > discussed on the mailing lists---I just returned from vacation and
+ > wasn't able to follow linux-kernel for a while).
+
+ Basically, the new setup is:
+
+ - we have "real address spaces" and "anonymous address spaces". The
+ difference is that an anonymous address space doesn't care about the
+ user-level page tables at all, so when we do a context switch into an
+ anonymous address space we just leave the previous address space
+ active.
+
+ The obvious use for a "anonymous address space" is any thread that
+ doesn't need any user mappings - all kernel threads basically fall into
+ this category, but even "real" threads can temporarily say that for
+ some amount of time they are not going to be interested in user space,
+ and that the scheduler might as well try to avoid wasting time on
+ switching the VM state around. Currently only the old-style bdflush
+ sync does that.
+
+ - "tsk->mm" points to the "real address space". For an anonymous process,
+ tsk->mm will be NULL, for the logical reason that an anonymous process
+ really doesn't _have_ a real address space at all.
+
+ - however, we obviously need to keep track of which address space we
+ "stole" for such an anonymous user. For that, we have "tsk->active_mm",
+ which shows what the currently active address space is.
+
+ The rule is that for a process with a real address space (ie tsk->mm is
+ non-NULL) the active_mm obviously always has to be the same as the real
+ one.
+
+ For a anonymous process, tsk->mm == NULL, and tsk->active_mm is the
+ "borrowed" mm while the anonymous process is running. When the
+ anonymous process gets scheduled away, the borrowed address space is
+ returned and cleared.
+
+ To support all that, the "struct mm_struct" now has two counters: a
+ "mm_users" counter that is how many "real address space users" there are,
+ and a "mm_count" counter that is the number of "lazy" users (ie anonymous
+ users) plus one if there are any real users.
+
+ Usually there is at least one real user, but it could be that the real
+ user exited on another CPU while a lazy user was still active, so you do
+ actually get cases where you have a address space that is _only_ used by
+ lazy users. That is often a short-lived state, because once that thread
+ gets scheduled away in favour of a real thread, the "zombie" mm gets
+ released because "mm_count" becomes zero.
+
+ Also, a new rule is that _nobody_ ever has "init_mm" as a real MM any
+ more. "init_mm" should be considered just a "lazy context when no other
+ context is available", and in fact it is mainly used just at bootup when
+ no real VM has yet been created. So code that used to check
+
+ if (current->mm == &init_mm)
+
+ should generally just do
+
+ if (!current->mm)
+
+ instead (which makes more sense anyway - the test is basically one of "do
+ we have a user context", and is generally done by the page fault handler
+ and things like that).
+
+ Anyway, I put a pre-patch-2.3.13-1 on ftp.kernel.org just a moment ago,
+ because it slightly changes the interfaces to accommodate the alpha (who
+ would have thought it, but the alpha actually ends up having one of the
+ ugliest context switch codes - unlike the other architectures where the MM
+ and register state is separate, the alpha PALcode joins the two, and you
+ need to switch both together).
+
+ (From http://marc.info/?l=linux-kernel&m=93337278602211&w=2)
diff --git a/Documentation/vm/arch_pgtable_helpers.rst b/Documentation/vm/arch_pgtable_helpers.rst
new file mode 100644
index 000000000..552567d86
--- /dev/null
+++ b/Documentation/vm/arch_pgtable_helpers.rst
@@ -0,0 +1,258 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+.. _arch_page_table_helpers:
+
+===============================
+Architecture Page Table Helpers
+===============================
+
+Generic MM expects architectures (with MMU) to provide helpers to create, access
+and modify page table entries at various level for different memory functions.
+These page table helpers need to conform to a common semantics across platforms.
+Following tables describe the expected semantics which can also be tested during
+boot via CONFIG_DEBUG_VM_PGTABLE option. All future changes in here or the debug
+test need to be in sync.
+
+======================
+PTE Page Table Helpers
+======================
+
++---------------------------+--------------------------------------------------+
+| pte_same | Tests whether both PTE entries are the same |
++---------------------------+--------------------------------------------------+
+| pte_bad | Tests a non-table mapped PTE |
++---------------------------+--------------------------------------------------+
+| pte_present | Tests a valid mapped PTE |
++---------------------------+--------------------------------------------------+
+| pte_young | Tests a young PTE |
++---------------------------+--------------------------------------------------+
+| pte_dirty | Tests a dirty PTE |
++---------------------------+--------------------------------------------------+
+| pte_write | Tests a writable PTE |
++---------------------------+--------------------------------------------------+
+| pte_special | Tests a special PTE |
++---------------------------+--------------------------------------------------+
+| pte_protnone | Tests a PROT_NONE PTE |
++---------------------------+--------------------------------------------------+
+| pte_devmap | Tests a ZONE_DEVICE mapped PTE |
++---------------------------+--------------------------------------------------+
+| pte_soft_dirty | Tests a soft dirty PTE |
++---------------------------+--------------------------------------------------+
+| pte_swp_soft_dirty | Tests a soft dirty swapped PTE |
++---------------------------+--------------------------------------------------+
+| pte_mkyoung | Creates a young PTE |
++---------------------------+--------------------------------------------------+
+| pte_mkold | Creates an old PTE |
++---------------------------+--------------------------------------------------+
+| pte_mkdirty | Creates a dirty PTE |
++---------------------------+--------------------------------------------------+
+| pte_mkclean | Creates a clean PTE |
++---------------------------+--------------------------------------------------+
+| pte_mkwrite | Creates a writable PTE |
++---------------------------+--------------------------------------------------+
+| pte_wrprotect | Creates a write protected PTE |
++---------------------------+--------------------------------------------------+
+| pte_mkspecial | Creates a special PTE |
++---------------------------+--------------------------------------------------+
+| pte_mkdevmap | Creates a ZONE_DEVICE mapped PTE |
++---------------------------+--------------------------------------------------+
+| pte_mksoft_dirty | Creates a soft dirty PTE |
++---------------------------+--------------------------------------------------+
+| pte_clear_soft_dirty | Clears a soft dirty PTE |
++---------------------------+--------------------------------------------------+
+| pte_swp_mksoft_dirty | Creates a soft dirty swapped PTE |
++---------------------------+--------------------------------------------------+
+| pte_swp_clear_soft_dirty | Clears a soft dirty swapped PTE |
++---------------------------+--------------------------------------------------+
+| pte_mknotpresent | Invalidates a mapped PTE |
++---------------------------+--------------------------------------------------+
+| ptep_get_and_clear | Clears a PTE |
++---------------------------+--------------------------------------------------+
+| ptep_get_and_clear_full | Clears a PTE |
++---------------------------+--------------------------------------------------+
+| ptep_test_and_clear_young | Clears young from a PTE |
++---------------------------+--------------------------------------------------+
+| ptep_set_wrprotect | Converts into a write protected PTE |
++---------------------------+--------------------------------------------------+
+| ptep_set_access_flags | Converts into a more permissive PTE |
++---------------------------+--------------------------------------------------+
+
+======================
+PMD Page Table Helpers
+======================
+
++---------------------------+--------------------------------------------------+
+| pmd_same | Tests whether both PMD entries are the same |
++---------------------------+--------------------------------------------------+
+| pmd_bad | Tests a non-table mapped PMD |
++---------------------------+--------------------------------------------------+
+| pmd_leaf | Tests a leaf mapped PMD |
++---------------------------+--------------------------------------------------+
+| pmd_huge | Tests a HugeTLB mapped PMD |
++---------------------------+--------------------------------------------------+
+| pmd_trans_huge | Tests a Transparent Huge Page (THP) at PMD |
++---------------------------+--------------------------------------------------+
+| pmd_present | Tests a valid mapped PMD |
++---------------------------+--------------------------------------------------+
+| pmd_young | Tests a young PMD |
++---------------------------+--------------------------------------------------+
+| pmd_dirty | Tests a dirty PMD |
++---------------------------+--------------------------------------------------+
+| pmd_write | Tests a writable PMD |
++---------------------------+--------------------------------------------------+
+| pmd_special | Tests a special PMD |
++---------------------------+--------------------------------------------------+
+| pmd_protnone | Tests a PROT_NONE PMD |
++---------------------------+--------------------------------------------------+
+| pmd_devmap | Tests a ZONE_DEVICE mapped PMD |
++---------------------------+--------------------------------------------------+
+| pmd_soft_dirty | Tests a soft dirty PMD |
++---------------------------+--------------------------------------------------+
+| pmd_swp_soft_dirty | Tests a soft dirty swapped PMD |
++---------------------------+--------------------------------------------------+
+| pmd_mkyoung | Creates a young PMD |
++---------------------------+--------------------------------------------------+
+| pmd_mkold | Creates an old PMD |
++---------------------------+--------------------------------------------------+
+| pmd_mkdirty | Creates a dirty PMD |
++---------------------------+--------------------------------------------------+
+| pmd_mkclean | Creates a clean PMD |
++---------------------------+--------------------------------------------------+
+| pmd_mkwrite | Creates a writable PMD |
++---------------------------+--------------------------------------------------+
+| pmd_wrprotect | Creates a write protected PMD |
++---------------------------+--------------------------------------------------+
+| pmd_mkspecial | Creates a special PMD |
++---------------------------+--------------------------------------------------+
+| pmd_mkdevmap | Creates a ZONE_DEVICE mapped PMD |
++---------------------------+--------------------------------------------------+
+| pmd_mksoft_dirty | Creates a soft dirty PMD |
++---------------------------+--------------------------------------------------+
+| pmd_clear_soft_dirty | Clears a soft dirty PMD |
++---------------------------+--------------------------------------------------+
+| pmd_swp_mksoft_dirty | Creates a soft dirty swapped PMD |
++---------------------------+--------------------------------------------------+
+| pmd_swp_clear_soft_dirty | Clears a soft dirty swapped PMD |
++---------------------------+--------------------------------------------------+
+| pmd_mkinvalid | Invalidates a mapped PMD [1] |
++---------------------------+--------------------------------------------------+
+| pmd_set_huge | Creates a PMD huge mapping |
++---------------------------+--------------------------------------------------+
+| pmd_clear_huge | Clears a PMD huge mapping |
++---------------------------+--------------------------------------------------+
+| pmdp_get_and_clear | Clears a PMD |
++---------------------------+--------------------------------------------------+
+| pmdp_get_and_clear_full | Clears a PMD |
++---------------------------+--------------------------------------------------+
+| pmdp_test_and_clear_young | Clears young from a PMD |
++---------------------------+--------------------------------------------------+
+| pmdp_set_wrprotect | Converts into a write protected PMD |
++---------------------------+--------------------------------------------------+
+| pmdp_set_access_flags | Converts into a more permissive PMD |
++---------------------------+--------------------------------------------------+
+
+======================
+PUD Page Table Helpers
+======================
+
++---------------------------+--------------------------------------------------+
+| pud_same | Tests whether both PUD entries are the same |
++---------------------------+--------------------------------------------------+
+| pud_bad | Tests a non-table mapped PUD |
++---------------------------+--------------------------------------------------+
+| pud_leaf | Tests a leaf mapped PUD |
++---------------------------+--------------------------------------------------+
+| pud_huge | Tests a HugeTLB mapped PUD |
++---------------------------+--------------------------------------------------+
+| pud_trans_huge | Tests a Transparent Huge Page (THP) at PUD |
++---------------------------+--------------------------------------------------+
+| pud_present | Tests a valid mapped PUD |
++---------------------------+--------------------------------------------------+
+| pud_young | Tests a young PUD |
++---------------------------+--------------------------------------------------+
+| pud_dirty | Tests a dirty PUD |
++---------------------------+--------------------------------------------------+
+| pud_write | Tests a writable PUD |
++---------------------------+--------------------------------------------------+
+| pud_devmap | Tests a ZONE_DEVICE mapped PUD |
++---------------------------+--------------------------------------------------+
+| pud_mkyoung | Creates a young PUD |
++---------------------------+--------------------------------------------------+
+| pud_mkold | Creates an old PUD |
++---------------------------+--------------------------------------------------+
+| pud_mkdirty | Creates a dirty PUD |
++---------------------------+--------------------------------------------------+
+| pud_mkclean | Creates a clean PUD |
++---------------------------+--------------------------------------------------+
+| pud_mkwrite | Creates a writable PUD |
++---------------------------+--------------------------------------------------+
+| pud_wrprotect | Creates a write protected PUD |
++---------------------------+--------------------------------------------------+
+| pud_mkdevmap | Creates a ZONE_DEVICE mapped PUD |
++---------------------------+--------------------------------------------------+
+| pud_mkinvalid | Invalidates a mapped PUD [1] |
++---------------------------+--------------------------------------------------+
+| pud_set_huge | Creates a PUD huge mapping |
++---------------------------+--------------------------------------------------+
+| pud_clear_huge | Clears a PUD huge mapping |
++---------------------------+--------------------------------------------------+
+| pudp_get_and_clear | Clears a PUD |
++---------------------------+--------------------------------------------------+
+| pudp_get_and_clear_full | Clears a PUD |
++---------------------------+--------------------------------------------------+
+| pudp_test_and_clear_young | Clears young from a PUD |
++---------------------------+--------------------------------------------------+
+| pudp_set_wrprotect | Converts into a write protected PUD |
++---------------------------+--------------------------------------------------+
+| pudp_set_access_flags | Converts into a more permissive PUD |
++---------------------------+--------------------------------------------------+
+
+==========================
+HugeTLB Page Table Helpers
+==========================
+
++---------------------------+--------------------------------------------------+
+| pte_huge | Tests a HugeTLB |
++---------------------------+--------------------------------------------------+
+| pte_mkhuge | Creates a HugeTLB |
++---------------------------+--------------------------------------------------+
+| huge_pte_dirty | Tests a dirty HugeTLB |
++---------------------------+--------------------------------------------------+
+| huge_pte_write | Tests a writable HugeTLB |
++---------------------------+--------------------------------------------------+
+| huge_pte_mkdirty | Creates a dirty HugeTLB |
++---------------------------+--------------------------------------------------+
+| huge_pte_mkwrite | Creates a writable HugeTLB |
++---------------------------+--------------------------------------------------+
+| huge_pte_wrprotect | Creates a write protected HugeTLB |
++---------------------------+--------------------------------------------------+
+| huge_ptep_get_and_clear | Clears a HugeTLB |
++---------------------------+--------------------------------------------------+
+| huge_ptep_set_wrprotect | Converts into a write protected HugeTLB |
++---------------------------+--------------------------------------------------+
+| huge_ptep_set_access_flags | Converts into a more permissive HugeTLB |
++---------------------------+--------------------------------------------------+
+
+========================
+SWAP Page Table Helpers
+========================
+
++---------------------------+--------------------------------------------------+
+| __pte_to_swp_entry | Creates a swapped entry (arch) from a mapped PTE |
++---------------------------+--------------------------------------------------+
+| __swp_to_pte_entry | Creates a mapped PTE from a swapped entry (arch) |
++---------------------------+--------------------------------------------------+
+| __pmd_to_swp_entry | Creates a swapped entry (arch) from a mapped PMD |
++---------------------------+--------------------------------------------------+
+| __swp_to_pmd_entry | Creates a mapped PMD from a swapped entry (arch) |
++---------------------------+--------------------------------------------------+
+| is_migration_entry | Tests a migration (read or write) swapped entry |
++---------------------------+--------------------------------------------------+
+| is_write_migration_entry | Tests a write migration swapped entry |
++---------------------------+--------------------------------------------------+
+| make_migration_entry_read | Converts into read migration swapped entry |
++---------------------------+--------------------------------------------------+
+| make_migration_entry | Creates a migration swapped entry (read or write)|
++---------------------------+--------------------------------------------------+
+
+[1] https://lore.kernel.org/linux-mm/20181017020930.GN30832@redhat.com/
diff --git a/Documentation/vm/balance.rst b/Documentation/vm/balance.rst
new file mode 100644
index 000000000..6a1fadf3e
--- /dev/null
+++ b/Documentation/vm/balance.rst
@@ -0,0 +1,102 @@
+.. _balance:
+
+================
+Memory Balancing
+================
+
+Started Jan 2000 by Kanoj Sarcar <kanoj@sgi.com>
+
+Memory balancing is needed for !__GFP_ATOMIC and !__GFP_KSWAPD_RECLAIM as
+well as for non __GFP_IO allocations.
+
+The first reason why a caller may avoid reclaim is that the caller can not
+sleep due to holding a spinlock or is in interrupt context. The second may
+be that the caller is willing to fail the allocation without incurring the
+overhead of page reclaim. This may happen for opportunistic high-order
+allocation requests that have order-0 fallback options. In such cases,
+the caller may also wish to avoid waking kswapd.
+
+__GFP_IO allocation requests are made to prevent file system deadlocks.
+
+In the absence of non sleepable allocation requests, it seems detrimental
+to be doing balancing. Page reclamation can be kicked off lazily, that
+is, only when needed (aka zone free memory is 0), instead of making it
+a proactive process.
+
+That being said, the kernel should try to fulfill requests for direct
+mapped pages from the direct mapped pool, instead of falling back on
+the dma pool, so as to keep the dma pool filled for dma requests (atomic
+or not). A similar argument applies to highmem and direct mapped pages.
+OTOH, if there is a lot of free dma pages, it is preferable to satisfy
+regular memory requests by allocating one from the dma pool, instead
+of incurring the overhead of regular zone balancing.
+
+In 2.2, memory balancing/page reclamation would kick off only when the
+_total_ number of free pages fell below 1/64 th of total memory. With the
+right ratio of dma and regular memory, it is quite possible that balancing
+would not be done even when the dma zone was completely empty. 2.2 has
+been running production machines of varying memory sizes, and seems to be
+doing fine even with the presence of this problem. In 2.3, due to
+HIGHMEM, this problem is aggravated.
+
+In 2.3, zone balancing can be done in one of two ways: depending on the
+zone size (and possibly of the size of lower class zones), we can decide
+at init time how many free pages we should aim for while balancing any
+zone. The good part is, while balancing, we do not need to look at sizes
+of lower class zones, the bad part is, we might do too frequent balancing
+due to ignoring possibly lower usage in the lower class zones. Also,
+with a slight change in the allocation routine, it is possible to reduce
+the memclass() macro to be a simple equality.
+
+Another possible solution is that we balance only when the free memory
+of a zone _and_ all its lower class zones falls below 1/64th of the
+total memory in the zone and its lower class zones. This fixes the 2.2
+balancing problem, and stays as close to 2.2 behavior as possible. Also,
+the balancing algorithm works the same way on the various architectures,
+which have different numbers and types of zones. If we wanted to get
+fancy, we could assign different weights to free pages in different
+zones in the future.
+
+Note that if the size of the regular zone is huge compared to dma zone,
+it becomes less significant to consider the free dma pages while
+deciding whether to balance the regular zone. The first solution
+becomes more attractive then.
+
+The appended patch implements the second solution. It also "fixes" two
+problems: first, kswapd is woken up as in 2.2 on low memory conditions
+for non-sleepable allocations. Second, the HIGHMEM zone is also balanced,
+so as to give a fighting chance for replace_with_highmem() to get a
+HIGHMEM page, as well as to ensure that HIGHMEM allocations do not
+fall back into regular zone. This also makes sure that HIGHMEM pages
+are not leaked (for example, in situations where a HIGHMEM page is in
+the swapcache but is not being used by anyone)
+
+kswapd also needs to know about the zones it should balance. kswapd is
+primarily needed in a situation where balancing can not be done,
+probably because all allocation requests are coming from intr context
+and all process contexts are sleeping. For 2.3, kswapd does not really
+need to balance the highmem zone, since intr context does not request
+highmem pages. kswapd looks at the zone_wake_kswapd field in the zone
+structure to decide whether a zone needs balancing.
+
+Page stealing from process memory and shm is done if stealing the page would
+alleviate memory pressure on any zone in the page's node that has fallen below
+its watermark.
+
+watemark[WMARK_MIN/WMARK_LOW/WMARK_HIGH]/low_on_memory/zone_wake_kswapd: These
+are per-zone fields, used to determine when a zone needs to be balanced. When
+the number of pages falls below watermark[WMARK_MIN], the hysteric field
+low_on_memory gets set. This stays set till the number of free pages becomes
+watermark[WMARK_HIGH]. When low_on_memory is set, page allocation requests will
+try to free some pages in the zone (providing GFP_WAIT is set in the request).
+Orthogonal to this, is the decision to poke kswapd to free some zone pages.
+That decision is not hysteresis based, and is done when the number of free
+pages is below watermark[WMARK_LOW]; in which case zone_wake_kswapd is also set.
+
+
+(Good) Ideas that I have heard:
+
+1. Dynamic experience should influence balancing: number of failed requests
+ for a zone can be tracked and fed into the balancing scheme (jalvo@mbay.net)
+2. Implement a replace_with_highmem()-like replace_with_regular() to preserve
+ dma pages. (lkd@tantalophile.demon.co.uk)
diff --git a/Documentation/vm/cleancache.rst b/Documentation/vm/cleancache.rst
new file mode 100644
index 000000000..68cba9131
--- /dev/null
+++ b/Documentation/vm/cleancache.rst
@@ -0,0 +1,296 @@
+.. _cleancache:
+
+==========
+Cleancache
+==========
+
+Motivation
+==========
+
+Cleancache is a new optional feature provided by the VFS layer that
+potentially dramatically increases page cache effectiveness for
+many workloads in many environments at a negligible cost.
+
+Cleancache can be thought of as a page-granularity victim cache for clean
+pages that the kernel's pageframe replacement algorithm (PFRA) would like
+to keep around, but can't since there isn't enough memory. So when the
+PFRA "evicts" a page, it first attempts to use cleancache code to
+put the data contained in that page into "transcendent memory", memory
+that is not directly accessible or addressable by the kernel and is
+of unknown and possibly time-varying size.
+
+Later, when a cleancache-enabled filesystem wishes to access a page
+in a file on disk, it first checks cleancache to see if it already
+contains it; if it does, the page of data is copied into the kernel
+and a disk access is avoided.
+
+Transcendent memory "drivers" for cleancache are currently implemented
+in Xen (using hypervisor memory) and zcache (using in-kernel compressed
+memory) and other implementations are in development.
+
+:ref:`FAQs <faq>` are included below.
+
+Implementation Overview
+=======================
+
+A cleancache "backend" that provides transcendent memory registers itself
+to the kernel's cleancache "frontend" by calling cleancache_register_ops,
+passing a pointer to a cleancache_ops structure with funcs set appropriately.
+The functions provided must conform to certain semantics as follows:
+
+Most important, cleancache is "ephemeral". Pages which are copied into
+cleancache have an indefinite lifetime which is completely unknowable
+by the kernel and so may or may not still be in cleancache at any later time.
+Thus, as its name implies, cleancache is not suitable for dirty pages.
+Cleancache has complete discretion over what pages to preserve and what
+pages to discard and when.
+
+Mounting a cleancache-enabled filesystem should call "init_fs" to obtain a
+pool id which, if positive, must be saved in the filesystem's superblock;
+a negative return value indicates failure. A "put_page" will copy a
+(presumably about-to-be-evicted) page into cleancache and associate it with
+the pool id, a file key, and a page index into the file. (The combination
+of a pool id, a file key, and an index is sometimes called a "handle".)
+A "get_page" will copy the page, if found, from cleancache into kernel memory.
+An "invalidate_page" will ensure the page no longer is present in cleancache;
+an "invalidate_inode" will invalidate all pages associated with the specified
+file; and, when a filesystem is unmounted, an "invalidate_fs" will invalidate
+all pages in all files specified by the given pool id and also surrender
+the pool id.
+
+An "init_shared_fs", like init_fs, obtains a pool id but tells cleancache
+to treat the pool as shared using a 128-bit UUID as a key. On systems
+that may run multiple kernels (such as hard partitioned or virtualized
+systems) that may share a clustered filesystem, and where cleancache
+may be shared among those kernels, calls to init_shared_fs that specify the
+same UUID will receive the same pool id, thus allowing the pages to
+be shared. Note that any security requirements must be imposed outside
+of the kernel (e.g. by "tools" that control cleancache). Or a
+cleancache implementation can simply disable shared_init by always
+returning a negative value.
+
+If a get_page is successful on a non-shared pool, the page is invalidated
+(thus making cleancache an "exclusive" cache). On a shared pool, the page
+is NOT invalidated on a successful get_page so that it remains accessible to
+other sharers. The kernel is responsible for ensuring coherency between
+cleancache (shared or not), the page cache, and the filesystem, using
+cleancache invalidate operations as required.
+
+Note that cleancache must enforce put-put-get coherency and get-get
+coherency. For the former, if two puts are made to the same handle but
+with different data, say AAA by the first put and BBB by the second, a
+subsequent get can never return the stale data (AAA). For get-get coherency,
+if a get for a given handle fails, subsequent gets for that handle will
+never succeed unless preceded by a successful put with that handle.
+
+Last, cleancache provides no SMP serialization guarantees; if two
+different Linux threads are simultaneously putting and invalidating a page
+with the same handle, the results are indeterminate. Callers must
+lock the page to ensure serial behavior.
+
+Cleancache Performance Metrics
+==============================
+
+If properly configured, monitoring of cleancache is done via debugfs in
+the `/sys/kernel/debug/cleancache` directory. The effectiveness of cleancache
+can be measured (across all filesystems) with:
+
+``succ_gets``
+ number of gets that were successful
+
+``failed_gets``
+ number of gets that failed
+
+``puts``
+ number of puts attempted (all "succeed")
+
+``invalidates``
+ number of invalidates attempted
+
+A backend implementation may provide additional metrics.
+
+.. _faq:
+
+FAQ
+===
+
+* Where's the value? (Andrew Morton)
+
+Cleancache provides a significant performance benefit to many workloads
+in many environments with negligible overhead by improving the
+effectiveness of the pagecache. Clean pagecache pages are
+saved in transcendent memory (RAM that is otherwise not directly
+addressable to the kernel); fetching those pages later avoids "refaults"
+and thus disk reads.
+
+Cleancache (and its sister code "frontswap") provide interfaces for
+this transcendent memory (aka "tmem"), which conceptually lies between
+fast kernel-directly-addressable RAM and slower DMA/asynchronous devices.
+Disallowing direct kernel or userland reads/writes to tmem
+is ideal when data is transformed to a different form and size (such
+as with compression) or secretly moved (as might be useful for write-
+balancing for some RAM-like devices). Evicted page-cache pages (and
+swap pages) are a great use for this kind of slower-than-RAM-but-much-
+faster-than-disk transcendent memory, and the cleancache (and frontswap)
+"page-object-oriented" specification provides a nice way to read and
+write -- and indirectly "name" -- the pages.
+
+In the virtual case, the whole point of virtualization is to statistically
+multiplex physical resources across the varying demands of multiple
+virtual machines. This is really hard to do with RAM and efforts to
+do it well with no kernel change have essentially failed (except in some
+well-publicized special-case workloads). Cleancache -- and frontswap --
+with a fairly small impact on the kernel, provide a huge amount
+of flexibility for more dynamic, flexible RAM multiplexing.
+Specifically, the Xen Transcendent Memory backend allows otherwise
+"fallow" hypervisor-owned RAM to not only be "time-shared" between multiple
+virtual machines, but the pages can be compressed and deduplicated to
+optimize RAM utilization. And when guest OS's are induced to surrender
+underutilized RAM (e.g. with "self-ballooning"), page cache pages
+are the first to go, and cleancache allows those pages to be
+saved and reclaimed if overall host system memory conditions allow.
+
+And the identical interface used for cleancache can be used in
+physical systems as well. The zcache driver acts as a memory-hungry
+device that stores pages of data in a compressed state. And
+the proposed "RAMster" driver shares RAM across multiple physical
+systems.
+
+* Why does cleancache have its sticky fingers so deep inside the
+ filesystems and VFS? (Andrew Morton and Christoph Hellwig)
+
+The core hooks for cleancache in VFS are in most cases a single line
+and the minimum set are placed precisely where needed to maintain
+coherency (via cleancache_invalidate operations) between cleancache,
+the page cache, and disk. All hooks compile into nothingness if
+cleancache is config'ed off and turn into a function-pointer-
+compare-to-NULL if config'ed on but no backend claims the ops
+functions, or to a compare-struct-element-to-negative if a
+backend claims the ops functions but a filesystem doesn't enable
+cleancache.
+
+Some filesystems are built entirely on top of VFS and the hooks
+in VFS are sufficient, so don't require an "init_fs" hook; the
+initial implementation of cleancache didn't provide this hook.
+But for some filesystems (such as btrfs), the VFS hooks are
+incomplete and one or more hooks in fs-specific code are required.
+And for some other filesystems, such as tmpfs, cleancache may
+be counterproductive. So it seemed prudent to require a filesystem
+to "opt in" to use cleancache, which requires adding a hook in
+each filesystem. Not all filesystems are supported by cleancache
+only because they haven't been tested. The existing set should
+be sufficient to validate the concept, the opt-in approach means
+that untested filesystems are not affected, and the hooks in the
+existing filesystems should make it very easy to add more
+filesystems in the future.
+
+The total impact of the hooks to existing fs and mm files is only
+about 40 lines added (not counting comments and blank lines).
+
+* Why not make cleancache asynchronous and batched so it can more
+ easily interface with real devices with DMA instead of copying each
+ individual page? (Minchan Kim)
+
+The one-page-at-a-time copy semantics simplifies the implementation
+on both the frontend and backend and also allows the backend to
+do fancy things on-the-fly like page compression and
+page deduplication. And since the data is "gone" (copied into/out
+of the pageframe) before the cleancache get/put call returns,
+a great deal of race conditions and potential coherency issues
+are avoided. While the interface seems odd for a "real device"
+or for real kernel-addressable RAM, it makes perfect sense for
+transcendent memory.
+
+* Why is non-shared cleancache "exclusive"? And where is the
+ page "invalidated" after a "get"? (Minchan Kim)
+
+The main reason is to free up space in transcendent memory and
+to avoid unnecessary cleancache_invalidate calls. If you want inclusive,
+the page can be "put" immediately following the "get". If
+put-after-get for inclusive becomes common, the interface could
+be easily extended to add a "get_no_invalidate" call.
+
+The invalidate is done by the cleancache backend implementation.
+
+* What's the performance impact?
+
+Performance analysis has been presented at OLS'09 and LCA'10.
+Briefly, performance gains can be significant on most workloads,
+especially when memory pressure is high (e.g. when RAM is
+overcommitted in a virtual workload); and because the hooks are
+invoked primarily in place of or in addition to a disk read/write,
+overhead is negligible even in worst case workloads. Basically
+cleancache replaces I/O with memory-copy-CPU-overhead; on older
+single-core systems with slow memory-copy speeds, cleancache
+has little value, but in newer multicore machines, especially
+consolidated/virtualized machines, it has great value.
+
+* How do I add cleancache support for filesystem X? (Boaz Harrash)
+
+Filesystems that are well-behaved and conform to certain
+restrictions can utilize cleancache simply by making a call to
+cleancache_init_fs at mount time. Unusual, misbehaving, or
+poorly layered filesystems must either add additional hooks
+and/or undergo extensive additional testing... or should just
+not enable the optional cleancache.
+
+Some points for a filesystem to consider:
+
+ - The FS should be block-device-based (e.g. a ram-based FS such
+ as tmpfs should not enable cleancache)
+ - To ensure coherency/correctness, the FS must ensure that all
+ file removal or truncation operations either go through VFS or
+ add hooks to do the equivalent cleancache "invalidate" operations
+ - To ensure coherency/correctness, either inode numbers must
+ be unique across the lifetime of the on-disk file OR the
+ FS must provide an "encode_fh" function.
+ - The FS must call the VFS superblock alloc and deactivate routines
+ or add hooks to do the equivalent cleancache calls done there.
+ - To maximize performance, all pages fetched from the FS should
+ go through the do_mpag_readpage routine or the FS should add
+ hooks to do the equivalent (cf. btrfs)
+ - Currently, the FS blocksize must be the same as PAGESIZE. This
+ is not an architectural restriction, but no backends currently
+ support anything different.
+ - A clustered FS should invoke the "shared_init_fs" cleancache
+ hook to get best performance for some backends.
+
+* Why not use the KVA of the inode as the key? (Christoph Hellwig)
+
+If cleancache would use the inode virtual address instead of
+inode/filehandle, the pool id could be eliminated. But, this
+won't work because cleancache retains pagecache data pages
+persistently even when the inode has been pruned from the
+inode unused list, and only invalidates the data page if the file
+gets removed/truncated. So if cleancache used the inode kva,
+there would be potential coherency issues if/when the inode
+kva is reused for a different file. Alternately, if cleancache
+invalidated the pages when the inode kva was freed, much of the value
+of cleancache would be lost because the cache of pages in cleanache
+is potentially much larger than the kernel pagecache and is most
+useful if the pages survive inode cache removal.
+
+* Why is a global variable required?
+
+The cleancache_enabled flag is checked in all of the frequently-used
+cleancache hooks. The alternative is a function call to check a static
+variable. Since cleancache is enabled dynamically at runtime, systems
+that don't enable cleancache would suffer thousands (possibly
+tens-of-thousands) of unnecessary function calls per second. So the
+global variable allows cleancache to be enabled by default at compile
+time, but have insignificant performance impact when cleancache remains
+disabled at runtime.
+
+* Does cleanache work with KVM?
+
+The memory model of KVM is sufficiently different that a cleancache
+backend may have less value for KVM. This remains to be tested,
+especially in an overcommitted system.
+
+* Does cleancache work in userspace? It sounds useful for
+ memory hungry caches like web browsers. (Jamie Lokier)
+
+No plans yet, though we agree it sounds useful, at least for
+apps that bypass the page cache (e.g. O_DIRECT).
+
+Last updated: Dan Magenheimer, April 13 2011
diff --git a/Documentation/vm/free_page_reporting.rst b/Documentation/vm/free_page_reporting.rst
new file mode 100644
index 000000000..8c05e62d8
--- /dev/null
+++ b/Documentation/vm/free_page_reporting.rst
@@ -0,0 +1,40 @@
+.. _free_page_reporting:
+
+=====================
+Free Page Reporting
+=====================
+
+Free page reporting is an API by which a device can register to receive
+lists of pages that are currently unused by the system. This is useful in
+the case of virtualization where a guest is then able to use this data to
+notify the hypervisor that it is no longer using certain pages in memory.
+
+For the driver, typically a balloon driver, to use of this functionality
+it will allocate and initialize a page_reporting_dev_info structure. The
+field within the structure it will populate is the "report" function
+pointer used to process the scatterlist. It must also guarantee that it can
+handle at least PAGE_REPORTING_CAPACITY worth of scatterlist entries per
+call to the function. A call to page_reporting_register will register the
+page reporting interface with the reporting framework assuming no other
+page reporting devices are already registered.
+
+Once registered the page reporting API will begin reporting batches of
+pages to the driver. The API will start reporting pages 2 seconds after
+the interface is registered and will continue to do so 2 seconds after any
+page of a sufficiently high order is freed.
+
+Pages reported will be stored in the scatterlist passed to the reporting
+function with the final entry having the end bit set in entry nent - 1.
+While pages are being processed by the report function they will not be
+accessible to the allocator. Once the report function has been completed
+the pages will be returned to the free area from which they were obtained.
+
+Prior to removing a driver that is making use of free page reporting it
+is necessary to call page_reporting_unregister to have the
+page_reporting_dev_info structure that is currently in use by free page
+reporting removed. Doing this will prevent further reports from being
+issued via the interface. If another driver or the same driver is
+registered it is possible for it to resume where the previous driver had
+left off in terms of reporting free pages.
+
+Alexander Duyck, Dec 04, 2019
diff --git a/Documentation/vm/frontswap.rst b/Documentation/vm/frontswap.rst
new file mode 100644
index 000000000..1979f430c
--- /dev/null
+++ b/Documentation/vm/frontswap.rst
@@ -0,0 +1,293 @@
+.. _frontswap:
+
+=========
+Frontswap
+=========
+
+Frontswap provides a "transcendent memory" interface for swap pages.
+In some environments, dramatic performance savings may be obtained because
+swapped pages are saved in RAM (or a RAM-like device) instead of a swap disk.
+
+(Note, frontswap -- and :ref:`cleancache` (merged at 3.0) -- are the "frontends"
+and the only necessary changes to the core kernel for transcendent memory;
+all other supporting code -- the "backends" -- is implemented as drivers.
+See the LWN.net article `Transcendent memory in a nutshell`_
+for a detailed overview of frontswap and related kernel parts)
+
+.. _Transcendent memory in a nutshell: https://lwn.net/Articles/454795/
+
+Frontswap is so named because it can be thought of as the opposite of
+a "backing" store for a swap device. The storage is assumed to be
+a synchronous concurrency-safe page-oriented "pseudo-RAM device" conforming
+to the requirements of transcendent memory (such as Xen's "tmem", or
+in-kernel compressed memory, aka "zcache", or future RAM-like devices);
+this pseudo-RAM device is not directly accessible or addressable by the
+kernel and is of unknown and possibly time-varying size. The driver
+links itself to frontswap by calling frontswap_register_ops to set the
+frontswap_ops funcs appropriately and the functions it provides must
+conform to certain policies as follows:
+
+An "init" prepares the device to receive frontswap pages associated
+with the specified swap device number (aka "type"). A "store" will
+copy the page to transcendent memory and associate it with the type and
+offset associated with the page. A "load" will copy the page, if found,
+from transcendent memory into kernel memory, but will NOT remove the page
+from transcendent memory. An "invalidate_page" will remove the page
+from transcendent memory and an "invalidate_area" will remove ALL pages
+associated with the swap type (e.g., like swapoff) and notify the "device"
+to refuse further stores with that swap type.
+
+Once a page is successfully stored, a matching load on the page will normally
+succeed. So when the kernel finds itself in a situation where it needs
+to swap out a page, it first attempts to use frontswap. If the store returns
+success, the data has been successfully saved to transcendent memory and
+a disk write and, if the data is later read back, a disk read are avoided.
+If a store returns failure, transcendent memory has rejected the data, and the
+page can be written to swap as usual.
+
+If a backend chooses, frontswap can be configured as a "writethrough
+cache" by calling frontswap_writethrough(). In this mode, the reduction
+in swap device writes is lost (and also a non-trivial performance advantage)
+in order to allow the backend to arbitrarily "reclaim" space used to
+store frontswap pages to more completely manage its memory usage.
+
+Note that if a page is stored and the page already exists in transcendent memory
+(a "duplicate" store), either the store succeeds and the data is overwritten,
+or the store fails AND the page is invalidated. This ensures stale data may
+never be obtained from frontswap.
+
+If properly configured, monitoring of frontswap is done via debugfs in
+the `/sys/kernel/debug/frontswap` directory. The effectiveness of
+frontswap can be measured (across all swap devices) with:
+
+``failed_stores``
+ how many store attempts have failed
+
+``loads``
+ how many loads were attempted (all should succeed)
+
+``succ_stores``
+ how many store attempts have succeeded
+
+``invalidates``
+ how many invalidates were attempted
+
+A backend implementation may provide additional metrics.
+
+FAQ
+===
+
+* Where's the value?
+
+When a workload starts swapping, performance falls through the floor.
+Frontswap significantly increases performance in many such workloads by
+providing a clean, dynamic interface to read and write swap pages to
+"transcendent memory" that is otherwise not directly addressable to the kernel.
+This interface is ideal when data is transformed to a different form
+and size (such as with compression) or secretly moved (as might be
+useful for write-balancing for some RAM-like devices). Swap pages (and
+evicted page-cache pages) are a great use for this kind of slower-than-RAM-
+but-much-faster-than-disk "pseudo-RAM device" and the frontswap (and
+cleancache) interface to transcendent memory provides a nice way to read
+and write -- and indirectly "name" -- the pages.
+
+Frontswap -- and cleancache -- with a fairly small impact on the kernel,
+provides a huge amount of flexibility for more dynamic, flexible RAM
+utilization in various system configurations:
+
+In the single kernel case, aka "zcache", pages are compressed and
+stored in local memory, thus increasing the total anonymous pages
+that can be safely kept in RAM. Zcache essentially trades off CPU
+cycles used in compression/decompression for better memory utilization.
+Benchmarks have shown little or no impact when memory pressure is
+low while providing a significant performance improvement (25%+)
+on some workloads under high memory pressure.
+
+"RAMster" builds on zcache by adding "peer-to-peer" transcendent memory
+support for clustered systems. Frontswap pages are locally compressed
+as in zcache, but then "remotified" to another system's RAM. This
+allows RAM to be dynamically load-balanced back-and-forth as needed,
+i.e. when system A is overcommitted, it can swap to system B, and
+vice versa. RAMster can also be configured as a memory server so
+many servers in a cluster can swap, dynamically as needed, to a single
+server configured with a large amount of RAM... without pre-configuring
+how much of the RAM is available for each of the clients!
+
+In the virtual case, the whole point of virtualization is to statistically
+multiplex physical resources across the varying demands of multiple
+virtual machines. This is really hard to do with RAM and efforts to do
+it well with no kernel changes have essentially failed (except in some
+well-publicized special-case workloads).
+Specifically, the Xen Transcendent Memory backend allows otherwise
+"fallow" hypervisor-owned RAM to not only be "time-shared" between multiple
+virtual machines, but the pages can be compressed and deduplicated to
+optimize RAM utilization. And when guest OS's are induced to surrender
+underutilized RAM (e.g. with "selfballooning"), sudden unexpected
+memory pressure may result in swapping; frontswap allows those pages
+to be swapped to and from hypervisor RAM (if overall host system memory
+conditions allow), thus mitigating the potentially awful performance impact
+of unplanned swapping.
+
+A KVM implementation is underway and has been RFC'ed to lkml. And,
+using frontswap, investigation is also underway on the use of NVM as
+a memory extension technology.
+
+* Sure there may be performance advantages in some situations, but
+ what's the space/time overhead of frontswap?
+
+If CONFIG_FRONTSWAP is disabled, every frontswap hook compiles into
+nothingness and the only overhead is a few extra bytes per swapon'ed
+swap device. If CONFIG_FRONTSWAP is enabled but no frontswap "backend"
+registers, there is one extra global variable compared to zero for
+every swap page read or written. If CONFIG_FRONTSWAP is enabled
+AND a frontswap backend registers AND the backend fails every "store"
+request (i.e. provides no memory despite claiming it might),
+CPU overhead is still negligible -- and since every frontswap fail
+precedes a swap page write-to-disk, the system is highly likely
+to be I/O bound and using a small fraction of a percent of a CPU
+will be irrelevant anyway.
+
+As for space, if CONFIG_FRONTSWAP is enabled AND a frontswap backend
+registers, one bit is allocated for every swap page for every swap
+device that is swapon'd. This is added to the EIGHT bits (which
+was sixteen until about 2.6.34) that the kernel already allocates
+for every swap page for every swap device that is swapon'd. (Hugh
+Dickins has observed that frontswap could probably steal one of
+the existing eight bits, but let's worry about that minor optimization
+later.) For very large swap disks (which are rare) on a standard
+4K pagesize, this is 1MB per 32GB swap.
+
+When swap pages are stored in transcendent memory instead of written
+out to disk, there is a side effect that this may create more memory
+pressure that can potentially outweigh the other advantages. A
+backend, such as zcache, must implement policies to carefully (but
+dynamically) manage memory limits to ensure this doesn't happen.
+
+* OK, how about a quick overview of what this frontswap patch does
+ in terms that a kernel hacker can grok?
+
+Let's assume that a frontswap "backend" has registered during
+kernel initialization; this registration indicates that this
+frontswap backend has access to some "memory" that is not directly
+accessible by the kernel. Exactly how much memory it provides is
+entirely dynamic and random.
+
+Whenever a swap-device is swapon'd frontswap_init() is called,
+passing the swap device number (aka "type") as a parameter.
+This notifies frontswap to expect attempts to "store" swap pages
+associated with that number.
+
+Whenever the swap subsystem is readying a page to write to a swap
+device (c.f swap_writepage()), frontswap_store is called. Frontswap
+consults with the frontswap backend and if the backend says it does NOT
+have room, frontswap_store returns -1 and the kernel swaps the page
+to the swap device as normal. Note that the response from the frontswap
+backend is unpredictable to the kernel; it may choose to never accept a
+page, it could accept every ninth page, or it might accept every
+page. But if the backend does accept a page, the data from the page
+has already been copied and associated with the type and offset,
+and the backend guarantees the persistence of the data. In this case,
+frontswap sets a bit in the "frontswap_map" for the swap device
+corresponding to the page offset on the swap device to which it would
+otherwise have written the data.
+
+When the swap subsystem needs to swap-in a page (swap_readpage()),
+it first calls frontswap_load() which checks the frontswap_map to
+see if the page was earlier accepted by the frontswap backend. If
+it was, the page of data is filled from the frontswap backend and
+the swap-in is complete. If not, the normal swap-in code is
+executed to obtain the page of data from the real swap device.
+
+So every time the frontswap backend accepts a page, a swap device read
+and (potentially) a swap device write are replaced by a "frontswap backend
+store" and (possibly) a "frontswap backend loads", which are presumably much
+faster.
+
+* Can't frontswap be configured as a "special" swap device that is
+ just higher priority than any real swap device (e.g. like zswap,
+ or maybe swap-over-nbd/NFS)?
+
+No. First, the existing swap subsystem doesn't allow for any kind of
+swap hierarchy. Perhaps it could be rewritten to accommodate a hierarchy,
+but this would require fairly drastic changes. Even if it were
+rewritten, the existing swap subsystem uses the block I/O layer which
+assumes a swap device is fixed size and any page in it is linearly
+addressable. Frontswap barely touches the existing swap subsystem,
+and works around the constraints of the block I/O subsystem to provide
+a great deal of flexibility and dynamicity.
+
+For example, the acceptance of any swap page by the frontswap backend is
+entirely unpredictable. This is critical to the definition of frontswap
+backends because it grants completely dynamic discretion to the
+backend. In zcache, one cannot know a priori how compressible a page is.
+"Poorly" compressible pages can be rejected, and "poorly" can itself be
+defined dynamically depending on current memory constraints.
+
+Further, frontswap is entirely synchronous whereas a real swap
+device is, by definition, asynchronous and uses block I/O. The
+block I/O layer is not only unnecessary, but may perform "optimizations"
+that are inappropriate for a RAM-oriented device including delaying
+the write of some pages for a significant amount of time. Synchrony is
+required to ensure the dynamicity of the backend and to avoid thorny race
+conditions that would unnecessarily and greatly complicate frontswap
+and/or the block I/O subsystem. That said, only the initial "store"
+and "load" operations need be synchronous. A separate asynchronous thread
+is free to manipulate the pages stored by frontswap. For example,
+the "remotification" thread in RAMster uses standard asynchronous
+kernel sockets to move compressed frontswap pages to a remote machine.
+Similarly, a KVM guest-side implementation could do in-guest compression
+and use "batched" hypercalls.
+
+In a virtualized environment, the dynamicity allows the hypervisor
+(or host OS) to do "intelligent overcommit". For example, it can
+choose to accept pages only until host-swapping might be imminent,
+then force guests to do their own swapping.
+
+There is a downside to the transcendent memory specifications for
+frontswap: Since any "store" might fail, there must always be a real
+slot on a real swap device to swap the page. Thus frontswap must be
+implemented as a "shadow" to every swapon'd device with the potential
+capability of holding every page that the swap device might have held
+and the possibility that it might hold no pages at all. This means
+that frontswap cannot contain more pages than the total of swapon'd
+swap devices. For example, if NO swap device is configured on some
+installation, frontswap is useless. Swapless portable devices
+can still use frontswap but a backend for such devices must configure
+some kind of "ghost" swap device and ensure that it is never used.
+
+* Why this weird definition about "duplicate stores"? If a page
+ has been previously successfully stored, can't it always be
+ successfully overwritten?
+
+Nearly always it can, but no, sometimes it cannot. Consider an example
+where data is compressed and the original 4K page has been compressed
+to 1K. Now an attempt is made to overwrite the page with data that
+is non-compressible and so would take the entire 4K. But the backend
+has no more space. In this case, the store must be rejected. Whenever
+frontswap rejects a store that would overwrite, it also must invalidate
+the old data and ensure that it is no longer accessible. Since the
+swap subsystem then writes the new data to the read swap device,
+this is the correct course of action to ensure coherency.
+
+* What is frontswap_shrink for?
+
+When the (non-frontswap) swap subsystem swaps out a page to a real
+swap device, that page is only taking up low-value pre-allocated disk
+space. But if frontswap has placed a page in transcendent memory, that
+page may be taking up valuable real estate. The frontswap_shrink
+routine allows code outside of the swap subsystem to force pages out
+of the memory managed by frontswap and back into kernel-addressable memory.
+For example, in RAMster, a "suction driver" thread will attempt
+to "repatriate" pages sent to a remote machine back to the local machine;
+this is driven using the frontswap_shrink mechanism when memory pressure
+subsides.
+
+* Why does the frontswap patch create the new include file swapfile.h?
+
+The frontswap code depends on some swap-subsystem-internal data
+structures that have, over the years, moved back and forth between
+static and global. This seemed a reasonable compromise: Define
+them as global but declare them in a new include file that isn't
+included by the large number of source files that include swap.h.
+
+Dan Magenheimer, last updated April 9, 2012
diff --git a/Documentation/vm/highmem.rst b/Documentation/vm/highmem.rst
new file mode 100644
index 000000000..0f69a9fec
--- /dev/null
+++ b/Documentation/vm/highmem.rst
@@ -0,0 +1,147 @@
+.. _highmem:
+
+====================
+High Memory Handling
+====================
+
+By: Peter Zijlstra <a.p.zijlstra@chello.nl>
+
+.. contents:: :local:
+
+What Is High Memory?
+====================
+
+High memory (highmem) is used when the size of physical memory approaches or
+exceeds the maximum size of virtual memory. At that point it becomes
+impossible for the kernel to keep all of the available physical memory mapped
+at all times. This means the kernel needs to start using temporary mappings of
+the pieces of physical memory that it wants to access.
+
+The part of (physical) memory not covered by a permanent mapping is what we
+refer to as 'highmem'. There are various architecture dependent constraints on
+where exactly that border lies.
+
+In the i386 arch, for example, we choose to map the kernel into every process's
+VM space so that we don't have to pay the full TLB invalidation costs for
+kernel entry/exit. This means the available virtual memory space (4GiB on
+i386) has to be divided between user and kernel space.
+
+The traditional split for architectures using this approach is 3:1, 3GiB for
+userspace and the top 1GiB for kernel space::
+
+ +--------+ 0xffffffff
+ | Kernel |
+ +--------+ 0xc0000000
+ | |
+ | User |
+ | |
+ +--------+ 0x00000000
+
+This means that the kernel can at most map 1GiB of physical memory at any one
+time, but because we need virtual address space for other things - including
+temporary maps to access the rest of the physical memory - the actual direct
+map will typically be less (usually around ~896MiB).
+
+Other architectures that have mm context tagged TLBs can have separate kernel
+and user maps. Some hardware (like some ARMs), however, have limited virtual
+space when they use mm context tags.
+
+
+Temporary Virtual Mappings
+==========================
+
+The kernel contains several ways of creating temporary mappings:
+
+* vmap(). This can be used to make a long duration mapping of multiple
+ physical pages into a contiguous virtual space. It needs global
+ synchronization to unmap.
+
+* kmap(). This permits a short duration mapping of a single page. It needs
+ global synchronization, but is amortized somewhat. It is also prone to
+ deadlocks when using in a nested fashion, and so it is not recommended for
+ new code.
+
+* kmap_atomic(). This permits a very short duration mapping of a single
+ page. Since the mapping is restricted to the CPU that issued it, it
+ performs well, but the issuing task is therefore required to stay on that
+ CPU until it has finished, lest some other task displace its mappings.
+
+ kmap_atomic() may also be used by interrupt contexts, since it is does not
+ sleep and the caller may not sleep until after kunmap_atomic() is called.
+
+ It may be assumed that k[un]map_atomic() won't fail.
+
+
+Using kmap_atomic
+=================
+
+When and where to use kmap_atomic() is straightforward. It is used when code
+wants to access the contents of a page that might be allocated from high memory
+(see __GFP_HIGHMEM), for example a page in the pagecache. The API has two
+functions, and they can be used in a manner similar to the following::
+
+ /* Find the page of interest. */
+ struct page *page = find_get_page(mapping, offset);
+
+ /* Gain access to the contents of that page. */
+ void *vaddr = kmap_atomic(page);
+
+ /* Do something to the contents of that page. */
+ memset(vaddr, 0, PAGE_SIZE);
+
+ /* Unmap that page. */
+ kunmap_atomic(vaddr);
+
+Note that the kunmap_atomic() call takes the result of the kmap_atomic() call
+not the argument.
+
+If you need to map two pages because you want to copy from one page to
+another you need to keep the kmap_atomic calls strictly nested, like::
+
+ vaddr1 = kmap_atomic(page1);
+ vaddr2 = kmap_atomic(page2);
+
+ memcpy(vaddr1, vaddr2, PAGE_SIZE);
+
+ kunmap_atomic(vaddr2);
+ kunmap_atomic(vaddr1);
+
+
+Cost of Temporary Mappings
+==========================
+
+The cost of creating temporary mappings can be quite high. The arch has to
+manipulate the kernel's page tables, the data TLB and/or the MMU's registers.
+
+If CONFIG_HIGHMEM is not set, then the kernel will try and create a mapping
+simply with a bit of arithmetic that will convert the page struct address into
+a pointer to the page contents rather than juggling mappings about. In such a
+case, the unmap operation may be a null operation.
+
+If CONFIG_MMU is not set, then there can be no temporary mappings and no
+highmem. In such a case, the arithmetic approach will also be used.
+
+
+i386 PAE
+========
+
+The i386 arch, under some circumstances, will permit you to stick up to 64GiB
+of RAM into your 32-bit machine. This has a number of consequences:
+
+* Linux needs a page-frame structure for each page in the system and the
+ pageframes need to live in the permanent mapping, which means:
+
+* you can have 896M/sizeof(struct page) page-frames at most; with struct
+ page being 32-bytes that would end up being something in the order of 112G
+ worth of pages; the kernel, however, needs to store more than just
+ page-frames in that memory...
+
+* PAE makes your page tables larger - which slows the system down as more
+ data has to be accessed to traverse in TLB fills and the like. One
+ advantage is that PAE has more PTE bits and can provide advanced features
+ like NX and PAT.
+
+The general recommendation is that you don't use more than 8GiB on a 32-bit
+machine - although more might work for you and your workload, you're pretty
+much on your own - don't expect kernel developers to really care much if things
+come apart.
diff --git a/Documentation/vm/hmm.rst b/Documentation/vm/hmm.rst
new file mode 100644
index 000000000..09e28507f
--- /dev/null
+++ b/Documentation/vm/hmm.rst
@@ -0,0 +1,435 @@
+.. _hmm:
+
+=====================================
+Heterogeneous Memory Management (HMM)
+=====================================
+
+Provide infrastructure and helpers to integrate non-conventional memory (device
+memory like GPU on board memory) into regular kernel path, with the cornerstone
+of this being specialized struct page for such memory (see sections 5 to 7 of
+this document).
+
+HMM also provides optional helpers for SVM (Share Virtual Memory), i.e.,
+allowing a device to transparently access program addresses coherently with
+the CPU meaning that any valid pointer on the CPU is also a valid pointer
+for the device. This is becoming mandatory to simplify the use of advanced
+heterogeneous computing where GPU, DSP, or FPGA are used to perform various
+computations on behalf of a process.
+
+This document is divided as follows: in the first section I expose the problems
+related to using device specific memory allocators. In the second section, I
+expose the hardware limitations that are inherent to many platforms. The third
+section gives an overview of the HMM design. The fourth section explains how
+CPU page-table mirroring works and the purpose of HMM in this context. The
+fifth section deals with how device memory is represented inside the kernel.
+Finally, the last section presents a new migration helper that allows
+leveraging the device DMA engine.
+
+.. contents:: :local:
+
+Problems of using a device specific memory allocator
+====================================================
+
+Devices with a large amount of on board memory (several gigabytes) like GPUs
+have historically managed their memory through dedicated driver specific APIs.
+This creates a disconnect between memory allocated and managed by a device
+driver and regular application memory (private anonymous, shared memory, or
+regular file backed memory). From here on I will refer to this aspect as split
+address space. I use shared address space to refer to the opposite situation:
+i.e., one in which any application memory region can be used by a device
+transparently.
+
+Split address space happens because devices can only access memory allocated
+through a device specific API. This implies that all memory objects in a program
+are not equal from the device point of view which complicates large programs
+that rely on a wide set of libraries.
+
+Concretely, this means that code that wants to leverage devices like GPUs needs
+to copy objects between generically allocated memory (malloc, mmap private, mmap
+share) and memory allocated through the device driver API (this still ends up
+with an mmap but of the device file).
+
+For flat data sets (array, grid, image, ...) this isn't too hard to achieve but
+for complex data sets (list, tree, ...) it's hard to get right. Duplicating a
+complex data set needs to re-map all the pointer relations between each of its
+elements. This is error prone and programs get harder to debug because of the
+duplicate data set and addresses.
+
+Split address space also means that libraries cannot transparently use data
+they are getting from the core program or another library and thus each library
+might have to duplicate its input data set using the device specific memory
+allocator. Large projects suffer from this and waste resources because of the
+various memory copies.
+
+Duplicating each library API to accept as input or output memory allocated by
+each device specific allocator is not a viable option. It would lead to a
+combinatorial explosion in the library entry points.
+
+Finally, with the advance of high level language constructs (in C++ but in
+other languages too) it is now possible for the compiler to leverage GPUs and
+other devices without programmer knowledge. Some compiler identified patterns
+are only do-able with a shared address space. It is also more reasonable to use
+a shared address space for all other patterns.
+
+
+I/O bus, device memory characteristics
+======================================
+
+I/O buses cripple shared address spaces due to a few limitations. Most I/O
+buses only allow basic memory access from device to main memory; even cache
+coherency is often optional. Access to device memory from a CPU is even more
+limited. More often than not, it is not cache coherent.
+
+If we only consider the PCIE bus, then a device can access main memory (often
+through an IOMMU) and be cache coherent with the CPUs. However, it only allows
+a limited set of atomic operations from the device on main memory. This is worse
+in the other direction: the CPU can only access a limited range of the device
+memory and cannot perform atomic operations on it. Thus device memory cannot
+be considered the same as regular memory from the kernel point of view.
+
+Another crippling factor is the limited bandwidth (~32GBytes/s with PCIE 4.0
+and 16 lanes). This is 33 times less than the fastest GPU memory (1 TBytes/s).
+The final limitation is latency. Access to main memory from the device has an
+order of magnitude higher latency than when the device accesses its own memory.
+
+Some platforms are developing new I/O buses or additions/modifications to PCIE
+to address some of these limitations (OpenCAPI, CCIX). They mainly allow
+two-way cache coherency between CPU and device and allow all atomic operations the
+architecture supports. Sadly, not all platforms are following this trend and
+some major architectures are left without hardware solutions to these problems.
+
+So for shared address space to make sense, not only must we allow devices to
+access any memory but we must also permit any memory to be migrated to device
+memory while the device is using it (blocking CPU access while it happens).
+
+
+Shared address space and migration
+==================================
+
+HMM intends to provide two main features. The first one is to share the address
+space by duplicating the CPU page table in the device page table so the same
+address points to the same physical memory for any valid main memory address in
+the process address space.
+
+To achieve this, HMM offers a set of helpers to populate the device page table
+while keeping track of CPU page table updates. Device page table updates are
+not as easy as CPU page table updates. To update the device page table, you must
+allocate a buffer (or use a pool of pre-allocated buffers) and write GPU
+specific commands in it to perform the update (unmap, cache invalidations, and
+flush, ...). This cannot be done through common code for all devices. Hence
+why HMM provides helpers to factor out everything that can be while leaving the
+hardware specific details to the device driver.
+
+The second mechanism HMM provides is a new kind of ZONE_DEVICE memory that
+allows allocating a struct page for each page of device memory. Those pages
+are special because the CPU cannot map them. However, they allow migrating
+main memory to device memory using existing migration mechanisms and everything
+looks like a page that is swapped out to disk from the CPU point of view. Using a
+struct page gives the easiest and cleanest integration with existing mm
+mechanisms. Here again, HMM only provides helpers, first to hotplug new ZONE_DEVICE
+memory for the device memory and second to perform migration. Policy decisions
+of what and when to migrate is left to the device driver.
+
+Note that any CPU access to a device page triggers a page fault and a migration
+back to main memory. For example, when a page backing a given CPU address A is
+migrated from a main memory page to a device page, then any CPU access to
+address A triggers a page fault and initiates a migration back to main memory.
+
+With these two features, HMM not only allows a device to mirror process address
+space and keeps both CPU and device page tables synchronized, but also
+leverages device memory by migrating the part of the data set that is actively being
+used by the device.
+
+
+Address space mirroring implementation and API
+==============================================
+
+Address space mirroring's main objective is to allow duplication of a range of
+CPU page table into a device page table; HMM helps keep both synchronized. A
+device driver that wants to mirror a process address space must start with the
+registration of a mmu_interval_notifier::
+
+ int mmu_interval_notifier_insert(struct mmu_interval_notifier *interval_sub,
+ struct mm_struct *mm, unsigned long start,
+ unsigned long length,
+ const struct mmu_interval_notifier_ops *ops);
+
+During the ops->invalidate() callback the device driver must perform the
+update action to the range (mark range read only, or fully unmap, etc.). The
+device must complete the update before the driver callback returns.
+
+When the device driver wants to populate a range of virtual addresses, it can
+use::
+
+ int hmm_range_fault(struct hmm_range *range);
+
+It will trigger a page fault on missing or read-only entries if write access is
+requested (see below). Page faults use the generic mm page fault code path just
+like a CPU page fault.
+
+Both functions copy CPU page table entries into their pfns array argument. Each
+entry in that array corresponds to an address in the virtual range. HMM
+provides a set of flags to help the driver identify special CPU page table
+entries.
+
+Locking within the sync_cpu_device_pagetables() callback is the most important
+aspect the driver must respect in order to keep things properly synchronized.
+The usage pattern is::
+
+ int driver_populate_range(...)
+ {
+ struct hmm_range range;
+ ...
+
+ range.notifier = &interval_sub;
+ range.start = ...;
+ range.end = ...;
+ range.hmm_pfns = ...;
+
+ if (!mmget_not_zero(interval_sub->notifier.mm))
+ return -EFAULT;
+
+ again:
+ range.notifier_seq = mmu_interval_read_begin(&interval_sub);
+ mmap_read_lock(mm);
+ ret = hmm_range_fault(&range);
+ if (ret) {
+ mmap_read_unlock(mm);
+ if (ret == -EBUSY)
+ goto again;
+ return ret;
+ }
+ mmap_read_unlock(mm);
+
+ take_lock(driver->update);
+ if (mmu_interval_read_retry(&ni, range.notifier_seq) {
+ release_lock(driver->update);
+ goto again;
+ }
+
+ /* Use pfns array content to update device page table,
+ * under the update lock */
+
+ release_lock(driver->update);
+ return 0;
+ }
+
+The driver->update lock is the same lock that the driver takes inside its
+invalidate() callback. That lock must be held before calling
+mmu_interval_read_retry() to avoid any race with a concurrent CPU page table
+update.
+
+Leverage default_flags and pfn_flags_mask
+=========================================
+
+The hmm_range struct has 2 fields, default_flags and pfn_flags_mask, that specify
+fault or snapshot policy for the whole range instead of having to set them
+for each entry in the pfns array.
+
+For instance if the device driver wants pages for a range with at least read
+permission, it sets::
+
+ range->default_flags = HMM_PFN_REQ_FAULT;
+ range->pfn_flags_mask = 0;
+
+and calls hmm_range_fault() as described above. This will fill fault all pages
+in the range with at least read permission.
+
+Now let's say the driver wants to do the same except for one page in the range for
+which it wants to have write permission. Now driver set::
+
+ range->default_flags = HMM_PFN_REQ_FAULT;
+ range->pfn_flags_mask = HMM_PFN_REQ_WRITE;
+ range->pfns[index_of_write] = HMM_PFN_REQ_WRITE;
+
+With this, HMM will fault in all pages with at least read (i.e., valid) and for the
+address == range->start + (index_of_write << PAGE_SHIFT) it will fault with
+write permission i.e., if the CPU pte does not have write permission set then HMM
+will call handle_mm_fault().
+
+After hmm_range_fault completes the flag bits are set to the current state of
+the page tables, ie HMM_PFN_VALID | HMM_PFN_WRITE will be set if the page is
+writable.
+
+
+Represent and manage device memory from core kernel point of view
+=================================================================
+
+Several different designs were tried to support device memory. The first one
+used a device specific data structure to keep information about migrated memory
+and HMM hooked itself in various places of mm code to handle any access to
+addresses that were backed by device memory. It turns out that this ended up
+replicating most of the fields of struct page and also needed many kernel code
+paths to be updated to understand this new kind of memory.
+
+Most kernel code paths never try to access the memory behind a page
+but only care about struct page contents. Because of this, HMM switched to
+directly using struct page for device memory which left most kernel code paths
+unaware of the difference. We only need to make sure that no one ever tries to
+map those pages from the CPU side.
+
+Migration to and from device memory
+===================================
+
+Because the CPU cannot access device memory directly, the device driver must
+use hardware DMA or device specific load/store instructions to migrate data.
+The migrate_vma_setup(), migrate_vma_pages(), and migrate_vma_finalize()
+functions are designed to make drivers easier to write and to centralize common
+code across drivers.
+
+Before migrating pages to device private memory, special device private
+``struct page`` need to be created. These will be used as special "swap"
+page table entries so that a CPU process will fault if it tries to access
+a page that has been migrated to device private memory.
+
+These can be allocated and freed with::
+
+ struct resource *res;
+ struct dev_pagemap pagemap;
+
+ res = request_free_mem_region(&iomem_resource, /* number of bytes */,
+ "name of driver resource");
+ pagemap.type = MEMORY_DEVICE_PRIVATE;
+ pagemap.range.start = res->start;
+ pagemap.range.end = res->end;
+ pagemap.nr_range = 1;
+ pagemap.ops = &device_devmem_ops;
+ memremap_pages(&pagemap, numa_node_id());
+
+ memunmap_pages(&pagemap);
+ release_mem_region(pagemap.range.start, range_len(&pagemap.range));
+
+There are also devm_request_free_mem_region(), devm_memremap_pages(),
+devm_memunmap_pages(), and devm_release_mem_region() when the resources can
+be tied to a ``struct device``.
+
+The overall migration steps are similar to migrating NUMA pages within system
+memory (see :ref:`Page migration <page_migration>`) but the steps are split
+between device driver specific code and shared common code:
+
+1. ``mmap_read_lock()``
+
+ The device driver has to pass a ``struct vm_area_struct`` to
+ migrate_vma_setup() so the mmap_read_lock() or mmap_write_lock() needs to
+ be held for the duration of the migration.
+
+2. ``migrate_vma_setup(struct migrate_vma *args)``
+
+ The device driver initializes the ``struct migrate_vma`` fields and passes
+ the pointer to migrate_vma_setup(). The ``args->flags`` field is used to
+ filter which source pages should be migrated. For example, setting
+ ``MIGRATE_VMA_SELECT_SYSTEM`` will only migrate system memory and
+ ``MIGRATE_VMA_SELECT_DEVICE_PRIVATE`` will only migrate pages residing in
+ device private memory. If the latter flag is set, the ``args->pgmap_owner``
+ field is used to identify device private pages owned by the driver. This
+ avoids trying to migrate device private pages residing in other devices.
+ Currently only anonymous private VMA ranges can be migrated to or from
+ system memory and device private memory.
+
+ One of the first steps migrate_vma_setup() does is to invalidate other
+ device's MMUs with the ``mmu_notifier_invalidate_range_start(()`` and
+ ``mmu_notifier_invalidate_range_end()`` calls around the page table
+ walks to fill in the ``args->src`` array with PFNs to be migrated.
+ The ``invalidate_range_start()`` callback is passed a
+ ``struct mmu_notifier_range`` with the ``event`` field set to
+ ``MMU_NOTIFY_MIGRATE`` and the ``migrate_pgmap_owner`` field set to
+ the ``args->pgmap_owner`` field passed to migrate_vma_setup(). This is
+ allows the device driver to skip the invalidation callback and only
+ invalidate device private MMU mappings that are actually migrating.
+ This is explained more in the next section.
+
+ While walking the page tables, a ``pte_none()`` or ``is_zero_pfn()``
+ entry results in a valid "zero" PFN stored in the ``args->src`` array.
+ This lets the driver allocate device private memory and clear it instead
+ of copying a page of zeros. Valid PTE entries to system memory or
+ device private struct pages will be locked with ``lock_page()``, isolated
+ from the LRU (if system memory since device private pages are not on
+ the LRU), unmapped from the process, and a special migration PTE is
+ inserted in place of the original PTE.
+ migrate_vma_setup() also clears the ``args->dst`` array.
+
+3. The device driver allocates destination pages and copies source pages to
+ destination pages.
+
+ The driver checks each ``src`` entry to see if the ``MIGRATE_PFN_MIGRATE``
+ bit is set and skips entries that are not migrating. The device driver
+ can also choose to skip migrating a page by not filling in the ``dst``
+ array for that page.
+
+ The driver then allocates either a device private struct page or a
+ system memory page, locks the page with ``lock_page()``, and fills in the
+ ``dst`` array entry with::
+
+ dst[i] = migrate_pfn(page_to_pfn(dpage)) | MIGRATE_PFN_LOCKED;
+
+ Now that the driver knows that this page is being migrated, it can
+ invalidate device private MMU mappings and copy device private memory
+ to system memory or another device private page. The core Linux kernel
+ handles CPU page table invalidations so the device driver only has to
+ invalidate its own MMU mappings.
+
+ The driver can use ``migrate_pfn_to_page(src[i])`` to get the
+ ``struct page`` of the source and either copy the source page to the
+ destination or clear the destination device private memory if the pointer
+ is ``NULL`` meaning the source page was not populated in system memory.
+
+4. ``migrate_vma_pages()``
+
+ This step is where the migration is actually "committed".
+
+ If the source page was a ``pte_none()`` or ``is_zero_pfn()`` page, this
+ is where the newly allocated page is inserted into the CPU's page table.
+ This can fail if a CPU thread faults on the same page. However, the page
+ table is locked and only one of the new pages will be inserted.
+ The device driver will see that the ``MIGRATE_PFN_MIGRATE`` bit is cleared
+ if it loses the race.
+
+ If the source page was locked, isolated, etc. the source ``struct page``
+ information is now copied to destination ``struct page`` finalizing the
+ migration on the CPU side.
+
+5. Device driver updates device MMU page tables for pages still migrating,
+ rolling back pages not migrating.
+
+ If the ``src`` entry still has ``MIGRATE_PFN_MIGRATE`` bit set, the device
+ driver can update the device MMU and set the write enable bit if the
+ ``MIGRATE_PFN_WRITE`` bit is set.
+
+6. ``migrate_vma_finalize()``
+
+ This step replaces the special migration page table entry with the new
+ page's page table entry and releases the reference to the source and
+ destination ``struct page``.
+
+7. ``mmap_read_unlock()``
+
+ The lock can now be released.
+
+Memory cgroup (memcg) and rss accounting
+========================================
+
+For now, device memory is accounted as any regular page in rss counters (either
+anonymous if device page is used for anonymous, file if device page is used for
+file backed page, or shmem if device page is used for shared memory). This is a
+deliberate choice to keep existing applications, that might start using device
+memory without knowing about it, running unimpacted.
+
+A drawback is that the OOM killer might kill an application using a lot of
+device memory and not a lot of regular system memory and thus not freeing much
+system memory. We want to gather more real world experience on how applications
+and system react under memory pressure in the presence of device memory before
+deciding to account device memory differently.
+
+
+Same decision was made for memory cgroup. Device memory pages are accounted
+against same memory cgroup a regular page would be accounted to. This does
+simplify migration to and from device memory. This also means that migration
+back from device memory to regular memory cannot fail because it would
+go above memory cgroup limit. We might revisit this choice latter on once we
+get more experience in how device memory is used and its impact on memory
+resource control.
+
+
+Note that device memory can never be pinned by a device driver nor through GUP
+and thus such memory is always free upon process exit. Or when last reference
+is dropped in case of shared memory or file backed memory.
diff --git a/Documentation/vm/hugetlbfs_reserv.rst b/Documentation/vm/hugetlbfs_reserv.rst
new file mode 100644
index 000000000..f143954e0
--- /dev/null
+++ b/Documentation/vm/hugetlbfs_reserv.rst
@@ -0,0 +1,596 @@
+.. _hugetlbfs_reserve:
+
+=====================
+Hugetlbfs Reservation
+=====================
+
+Overview
+========
+
+Huge pages as described at :ref:`hugetlbpage` are typically
+preallocated for application use. These huge pages are instantiated in a
+task's address space at page fault time if the VMA indicates huge pages are
+to be used. If no huge page exists at page fault time, the task is sent
+a SIGBUS and often dies an unhappy death. Shortly after huge page support
+was added, it was determined that it would be better to detect a shortage
+of huge pages at mmap() time. The idea is that if there were not enough
+huge pages to cover the mapping, the mmap() would fail. This was first
+done with a simple check in the code at mmap() time to determine if there
+were enough free huge pages to cover the mapping. Like most things in the
+kernel, the code has evolved over time. However, the basic idea was to
+'reserve' huge pages at mmap() time to ensure that huge pages would be
+available for page faults in that mapping. The description below attempts to
+describe how huge page reserve processing is done in the v4.10 kernel.
+
+
+Audience
+========
+This description is primarily targeted at kernel developers who are modifying
+hugetlbfs code.
+
+
+The Data Structures
+===================
+
+resv_huge_pages
+ This is a global (per-hstate) count of reserved huge pages. Reserved
+ huge pages are only available to the task which reserved them.
+ Therefore, the number of huge pages generally available is computed
+ as (``free_huge_pages - resv_huge_pages``).
+Reserve Map
+ A reserve map is described by the structure::
+
+ struct resv_map {
+ struct kref refs;
+ spinlock_t lock;
+ struct list_head regions;
+ long adds_in_progress;
+ struct list_head region_cache;
+ long region_cache_count;
+ };
+
+ There is one reserve map for each huge page mapping in the system.
+ The regions list within the resv_map describes the regions within
+ the mapping. A region is described as::
+
+ struct file_region {
+ struct list_head link;
+ long from;
+ long to;
+ };
+
+ The 'from' and 'to' fields of the file region structure are huge page
+ indices into the mapping. Depending on the type of mapping, a
+ region in the reserv_map may indicate reservations exist for the
+ range, or reservations do not exist.
+Flags for MAP_PRIVATE Reservations
+ These are stored in the bottom bits of the reservation map pointer.
+
+ ``#define HPAGE_RESV_OWNER (1UL << 0)``
+ Indicates this task is the owner of the reservations
+ associated with the mapping.
+ ``#define HPAGE_RESV_UNMAPPED (1UL << 1)``
+ Indicates task originally mapping this range (and creating
+ reserves) has unmapped a page from this task (the child)
+ due to a failed COW.
+Page Flags
+ The PagePrivate page flag is used to indicate that a huge page
+ reservation must be restored when the huge page is freed. More
+ details will be discussed in the "Freeing huge pages" section.
+
+
+Reservation Map Location (Private or Shared)
+============================================
+
+A huge page mapping or segment is either private or shared. If private,
+it is typically only available to a single address space (task). If shared,
+it can be mapped into multiple address spaces (tasks). The location and
+semantics of the reservation map is significantly different for the two types
+of mappings. Location differences are:
+
+- For private mappings, the reservation map hangs off the VMA structure.
+ Specifically, vma->vm_private_data. This reserve map is created at the
+ time the mapping (mmap(MAP_PRIVATE)) is created.
+- For shared mappings, the reservation map hangs off the inode. Specifically,
+ inode->i_mapping->private_data. Since shared mappings are always backed
+ by files in the hugetlbfs filesystem, the hugetlbfs code ensures each inode
+ contains a reservation map. As a result, the reservation map is allocated
+ when the inode is created.
+
+
+Creating Reservations
+=====================
+Reservations are created when a huge page backed shared memory segment is
+created (shmget(SHM_HUGETLB)) or a mapping is created via mmap(MAP_HUGETLB).
+These operations result in a call to the routine hugetlb_reserve_pages()::
+
+ int hugetlb_reserve_pages(struct inode *inode,
+ long from, long to,
+ struct vm_area_struct *vma,
+ vm_flags_t vm_flags)
+
+The first thing hugetlb_reserve_pages() does is check if the NORESERVE
+flag was specified in either the shmget() or mmap() call. If NORESERVE
+was specified, then this routine returns immediately as no reservations
+are desired.
+
+The arguments 'from' and 'to' are huge page indices into the mapping or
+underlying file. For shmget(), 'from' is always 0 and 'to' corresponds to
+the length of the segment/mapping. For mmap(), the offset argument could
+be used to specify the offset into the underlying file. In such a case,
+the 'from' and 'to' arguments have been adjusted by this offset.
+
+One of the big differences between PRIVATE and SHARED mappings is the way
+in which reservations are represented in the reservation map.
+
+- For shared mappings, an entry in the reservation map indicates a reservation
+ exists or did exist for the corresponding page. As reservations are
+ consumed, the reservation map is not modified.
+- For private mappings, the lack of an entry in the reservation map indicates
+ a reservation exists for the corresponding page. As reservations are
+ consumed, entries are added to the reservation map. Therefore, the
+ reservation map can also be used to determine which reservations have
+ been consumed.
+
+For private mappings, hugetlb_reserve_pages() creates the reservation map and
+hangs it off the VMA structure. In addition, the HPAGE_RESV_OWNER flag is set
+to indicate this VMA owns the reservations.
+
+The reservation map is consulted to determine how many huge page reservations
+are needed for the current mapping/segment. For private mappings, this is
+always the value (to - from). However, for shared mappings it is possible that
+some reservations may already exist within the range (to - from). See the
+section :ref:`Reservation Map Modifications <resv_map_modifications>`
+for details on how this is accomplished.
+
+The mapping may be associated with a subpool. If so, the subpool is consulted
+to ensure there is sufficient space for the mapping. It is possible that the
+subpool has set aside reservations that can be used for the mapping. See the
+section :ref:`Subpool Reservations <sub_pool_resv>` for more details.
+
+After consulting the reservation map and subpool, the number of needed new
+reservations is known. The routine hugetlb_acct_memory() is called to check
+for and take the requested number of reservations. hugetlb_acct_memory()
+calls into routines that potentially allocate and adjust surplus page counts.
+However, within those routines the code is simply checking to ensure there
+are enough free huge pages to accommodate the reservation. If there are,
+the global reservation count resv_huge_pages is adjusted something like the
+following::
+
+ if (resv_needed <= (resv_huge_pages - free_huge_pages))
+ resv_huge_pages += resv_needed;
+
+Note that the global lock hugetlb_lock is held when checking and adjusting
+these counters.
+
+If there were enough free huge pages and the global count resv_huge_pages
+was adjusted, then the reservation map associated with the mapping is
+modified to reflect the reservations. In the case of a shared mapping, a
+file_region will exist that includes the range 'from' - 'to'. For private
+mappings, no modifications are made to the reservation map as lack of an
+entry indicates a reservation exists.
+
+If hugetlb_reserve_pages() was successful, the global reservation count and
+reservation map associated with the mapping will be modified as required to
+ensure reservations exist for the range 'from' - 'to'.
+
+.. _consume_resv:
+
+Consuming Reservations/Allocating a Huge Page
+=============================================
+
+Reservations are consumed when huge pages associated with the reservations
+are allocated and instantiated in the corresponding mapping. The allocation
+is performed within the routine alloc_huge_page()::
+
+ struct page *alloc_huge_page(struct vm_area_struct *vma,
+ unsigned long addr, int avoid_reserve)
+
+alloc_huge_page is passed a VMA pointer and a virtual address, so it can
+consult the reservation map to determine if a reservation exists. In addition,
+alloc_huge_page takes the argument avoid_reserve which indicates reserves
+should not be used even if it appears they have been set aside for the
+specified address. The avoid_reserve argument is most often used in the case
+of Copy on Write and Page Migration where additional copies of an existing
+page are being allocated.
+
+The helper routine vma_needs_reservation() is called to determine if a
+reservation exists for the address within the mapping(vma). See the section
+:ref:`Reservation Map Helper Routines <resv_map_helpers>` for detailed
+information on what this routine does.
+The value returned from vma_needs_reservation() is generally
+0 or 1. 0 if a reservation exists for the address, 1 if no reservation exists.
+If a reservation does not exist, and there is a subpool associated with the
+mapping the subpool is consulted to determine if it contains reservations.
+If the subpool contains reservations, one can be used for this allocation.
+However, in every case the avoid_reserve argument overrides the use of
+a reservation for the allocation. After determining whether a reservation
+exists and can be used for the allocation, the routine dequeue_huge_page_vma()
+is called. This routine takes two arguments related to reservations:
+
+- avoid_reserve, this is the same value/argument passed to alloc_huge_page()
+- chg, even though this argument is of type long only the values 0 or 1 are
+ passed to dequeue_huge_page_vma. If the value is 0, it indicates a
+ reservation exists (see the section "Memory Policy and Reservations" for
+ possible issues). If the value is 1, it indicates a reservation does not
+ exist and the page must be taken from the global free pool if possible.
+
+The free lists associated with the memory policy of the VMA are searched for
+a free page. If a page is found, the value free_huge_pages is decremented
+when the page is removed from the free list. If there was a reservation
+associated with the page, the following adjustments are made::
+
+ SetPagePrivate(page); /* Indicates allocating this page consumed
+ * a reservation, and if an error is
+ * encountered such that the page must be
+ * freed, the reservation will be restored. */
+ resv_huge_pages--; /* Decrement the global reservation count */
+
+Note, if no huge page can be found that satisfies the VMA's memory policy
+an attempt will be made to allocate one using the buddy allocator. This
+brings up the issue of surplus huge pages and overcommit which is beyond
+the scope reservations. Even if a surplus page is allocated, the same
+reservation based adjustments as above will be made: SetPagePrivate(page) and
+resv_huge_pages--.
+
+After obtaining a new huge page, (page)->private is set to the value of
+the subpool associated with the page if it exists. This will be used for
+subpool accounting when the page is freed.
+
+The routine vma_commit_reservation() is then called to adjust the reserve
+map based on the consumption of the reservation. In general, this involves
+ensuring the page is represented within a file_region structure of the region
+map. For shared mappings where the reservation was present, an entry
+in the reserve map already existed so no change is made. However, if there
+was no reservation in a shared mapping or this was a private mapping a new
+entry must be created.
+
+It is possible that the reserve map could have been changed between the call
+to vma_needs_reservation() at the beginning of alloc_huge_page() and the
+call to vma_commit_reservation() after the page was allocated. This would
+be possible if hugetlb_reserve_pages was called for the same page in a shared
+mapping. In such cases, the reservation count and subpool free page count
+will be off by one. This rare condition can be identified by comparing the
+return value from vma_needs_reservation and vma_commit_reservation. If such
+a race is detected, the subpool and global reserve counts are adjusted to
+compensate. See the section
+:ref:`Reservation Map Helper Routines <resv_map_helpers>` for more
+information on these routines.
+
+
+Instantiate Huge Pages
+======================
+
+After huge page allocation, the page is typically added to the page tables
+of the allocating task. Before this, pages in a shared mapping are added
+to the page cache and pages in private mappings are added to an anonymous
+reverse mapping. In both cases, the PagePrivate flag is cleared. Therefore,
+when a huge page that has been instantiated is freed no adjustment is made
+to the global reservation count (resv_huge_pages).
+
+
+Freeing Huge Pages
+==================
+
+Huge page freeing is performed by the routine free_huge_page(). This routine
+is the destructor for hugetlbfs compound pages. As a result, it is only
+passed a pointer to the page struct. When a huge page is freed, reservation
+accounting may need to be performed. This would be the case if the page was
+associated with a subpool that contained reserves, or the page is being freed
+on an error path where a global reserve count must be restored.
+
+The page->private field points to any subpool associated with the page.
+If the PagePrivate flag is set, it indicates the global reserve count should
+be adjusted (see the section
+:ref:`Consuming Reservations/Allocating a Huge Page <consume_resv>`
+for information on how these are set).
+
+The routine first calls hugepage_subpool_put_pages() for the page. If this
+routine returns a value of 0 (which does not equal the value passed 1) it
+indicates reserves are associated with the subpool, and this newly free page
+must be used to keep the number of subpool reserves above the minimum size.
+Therefore, the global resv_huge_pages counter is incremented in this case.
+
+If the PagePrivate flag was set in the page, the global resv_huge_pages counter
+will always be incremented.
+
+.. _sub_pool_resv:
+
+Subpool Reservations
+====================
+
+There is a struct hstate associated with each huge page size. The hstate
+tracks all huge pages of the specified size. A subpool represents a subset
+of pages within a hstate that is associated with a mounted hugetlbfs
+filesystem.
+
+When a hugetlbfs filesystem is mounted a min_size option can be specified
+which indicates the minimum number of huge pages required by the filesystem.
+If this option is specified, the number of huge pages corresponding to
+min_size are reserved for use by the filesystem. This number is tracked in
+the min_hpages field of a struct hugepage_subpool. At mount time,
+hugetlb_acct_memory(min_hpages) is called to reserve the specified number of
+huge pages. If they can not be reserved, the mount fails.
+
+The routines hugepage_subpool_get/put_pages() are called when pages are
+obtained from or released back to a subpool. They perform all subpool
+accounting, and track any reservations associated with the subpool.
+hugepage_subpool_get/put_pages are passed the number of huge pages by which
+to adjust the subpool 'used page' count (down for get, up for put). Normally,
+they return the same value that was passed or an error if not enough pages
+exist in the subpool.
+
+However, if reserves are associated with the subpool a return value less
+than the passed value may be returned. This return value indicates the
+number of additional global pool adjustments which must be made. For example,
+suppose a subpool contains 3 reserved huge pages and someone asks for 5.
+The 3 reserved pages associated with the subpool can be used to satisfy part
+of the request. But, 2 pages must be obtained from the global pools. To
+relay this information to the caller, the value 2 is returned. The caller
+is then responsible for attempting to obtain the additional two pages from
+the global pools.
+
+
+COW and Reservations
+====================
+
+Since shared mappings all point to and use the same underlying pages, the
+biggest reservation concern for COW is private mappings. In this case,
+two tasks can be pointing at the same previously allocated page. One task
+attempts to write to the page, so a new page must be allocated so that each
+task points to its own page.
+
+When the page was originally allocated, the reservation for that page was
+consumed. When an attempt to allocate a new page is made as a result of
+COW, it is possible that no free huge pages are free and the allocation
+will fail.
+
+When the private mapping was originally created, the owner of the mapping
+was noted by setting the HPAGE_RESV_OWNER bit in the pointer to the reservation
+map of the owner. Since the owner created the mapping, the owner owns all
+the reservations associated with the mapping. Therefore, when a write fault
+occurs and there is no page available, different action is taken for the owner
+and non-owner of the reservation.
+
+In the case where the faulting task is not the owner, the fault will fail and
+the task will typically receive a SIGBUS.
+
+If the owner is the faulting task, we want it to succeed since it owned the
+original reservation. To accomplish this, the page is unmapped from the
+non-owning task. In this way, the only reference is from the owning task.
+In addition, the HPAGE_RESV_UNMAPPED bit is set in the reservation map pointer
+of the non-owning task. The non-owning task may receive a SIGBUS if it later
+faults on a non-present page. But, the original owner of the
+mapping/reservation will behave as expected.
+
+
+.. _resv_map_modifications:
+
+Reservation Map Modifications
+=============================
+
+The following low level routines are used to make modifications to a
+reservation map. Typically, these routines are not called directly. Rather,
+a reservation map helper routine is called which calls one of these low level
+routines. These low level routines are fairly well documented in the source
+code (mm/hugetlb.c). These routines are::
+
+ long region_chg(struct resv_map *resv, long f, long t);
+ long region_add(struct resv_map *resv, long f, long t);
+ void region_abort(struct resv_map *resv, long f, long t);
+ long region_count(struct resv_map *resv, long f, long t);
+
+Operations on the reservation map typically involve two operations:
+
+1) region_chg() is called to examine the reserve map and determine how
+ many pages in the specified range [f, t) are NOT currently represented.
+
+ The calling code performs global checks and allocations to determine if
+ there are enough huge pages for the operation to succeed.
+
+2)
+ a) If the operation can succeed, region_add() is called to actually modify
+ the reservation map for the same range [f, t) previously passed to
+ region_chg().
+ b) If the operation can not succeed, region_abort is called for the same
+ range [f, t) to abort the operation.
+
+Note that this is a two step process where region_add() and region_abort()
+are guaranteed to succeed after a prior call to region_chg() for the same
+range. region_chg() is responsible for pre-allocating any data structures
+necessary to ensure the subsequent operations (specifically region_add()))
+will succeed.
+
+As mentioned above, region_chg() determines the number of pages in the range
+which are NOT currently represented in the map. This number is returned to
+the caller. region_add() returns the number of pages in the range added to
+the map. In most cases, the return value of region_add() is the same as the
+return value of region_chg(). However, in the case of shared mappings it is
+possible for changes to the reservation map to be made between the calls to
+region_chg() and region_add(). In this case, the return value of region_add()
+will not match the return value of region_chg(). It is likely that in such
+cases global counts and subpool accounting will be incorrect and in need of
+adjustment. It is the responsibility of the caller to check for this condition
+and make the appropriate adjustments.
+
+The routine region_del() is called to remove regions from a reservation map.
+It is typically called in the following situations:
+
+- When a file in the hugetlbfs filesystem is being removed, the inode will
+ be released and the reservation map freed. Before freeing the reservation
+ map, all the individual file_region structures must be freed. In this case
+ region_del is passed the range [0, LONG_MAX).
+- When a hugetlbfs file is being truncated. In this case, all allocated pages
+ after the new file size must be freed. In addition, any file_region entries
+ in the reservation map past the new end of file must be deleted. In this
+ case, region_del is passed the range [new_end_of_file, LONG_MAX).
+- When a hole is being punched in a hugetlbfs file. In this case, huge pages
+ are removed from the middle of the file one at a time. As the pages are
+ removed, region_del() is called to remove the corresponding entry from the
+ reservation map. In this case, region_del is passed the range
+ [page_idx, page_idx + 1).
+
+In every case, region_del() will return the number of pages removed from the
+reservation map. In VERY rare cases, region_del() can fail. This can only
+happen in the hole punch case where it has to split an existing file_region
+entry and can not allocate a new structure. In this error case, region_del()
+will return -ENOMEM. The problem here is that the reservation map will
+indicate that there is a reservation for the page. However, the subpool and
+global reservation counts will not reflect the reservation. To handle this
+situation, the routine hugetlb_fix_reserve_counts() is called to adjust the
+counters so that they correspond with the reservation map entry that could
+not be deleted.
+
+region_count() is called when unmapping a private huge page mapping. In
+private mappings, the lack of a entry in the reservation map indicates that
+a reservation exists. Therefore, by counting the number of entries in the
+reservation map we know how many reservations were consumed and how many are
+outstanding (outstanding = (end - start) - region_count(resv, start, end)).
+Since the mapping is going away, the subpool and global reservation counts
+are decremented by the number of outstanding reservations.
+
+.. _resv_map_helpers:
+
+Reservation Map Helper Routines
+===============================
+
+Several helper routines exist to query and modify the reservation maps.
+These routines are only interested with reservations for a specific huge
+page, so they just pass in an address instead of a range. In addition,
+they pass in the associated VMA. From the VMA, the type of mapping (private
+or shared) and the location of the reservation map (inode or VMA) can be
+determined. These routines simply call the underlying routines described
+in the section "Reservation Map Modifications". However, they do take into
+account the 'opposite' meaning of reservation map entries for private and
+shared mappings and hide this detail from the caller::
+
+ long vma_needs_reservation(struct hstate *h,
+ struct vm_area_struct *vma,
+ unsigned long addr)
+
+This routine calls region_chg() for the specified page. If no reservation
+exists, 1 is returned. If a reservation exists, 0 is returned::
+
+ long vma_commit_reservation(struct hstate *h,
+ struct vm_area_struct *vma,
+ unsigned long addr)
+
+This calls region_add() for the specified page. As in the case of region_chg
+and region_add, this routine is to be called after a previous call to
+vma_needs_reservation. It will add a reservation entry for the page. It
+returns 1 if the reservation was added and 0 if not. The return value should
+be compared with the return value of the previous call to
+vma_needs_reservation. An unexpected difference indicates the reservation
+map was modified between calls::
+
+ void vma_end_reservation(struct hstate *h,
+ struct vm_area_struct *vma,
+ unsigned long addr)
+
+This calls region_abort() for the specified page. As in the case of region_chg
+and region_abort, this routine is to be called after a previous call to
+vma_needs_reservation. It will abort/end the in progress reservation add
+operation::
+
+ long vma_add_reservation(struct hstate *h,
+ struct vm_area_struct *vma,
+ unsigned long addr)
+
+This is a special wrapper routine to help facilitate reservation cleanup
+on error paths. It is only called from the routine restore_reserve_on_error().
+This routine is used in conjunction with vma_needs_reservation in an attempt
+to add a reservation to the reservation map. It takes into account the
+different reservation map semantics for private and shared mappings. Hence,
+region_add is called for shared mappings (as an entry present in the map
+indicates a reservation), and region_del is called for private mappings (as
+the absence of an entry in the map indicates a reservation). See the section
+"Reservation cleanup in error paths" for more information on what needs to
+be done on error paths.
+
+
+Reservation Cleanup in Error Paths
+==================================
+
+As mentioned in the section
+:ref:`Reservation Map Helper Routines <resv_map_helpers>`, reservation
+map modifications are performed in two steps. First vma_needs_reservation
+is called before a page is allocated. If the allocation is successful,
+then vma_commit_reservation is called. If not, vma_end_reservation is called.
+Global and subpool reservation counts are adjusted based on success or failure
+of the operation and all is well.
+
+Additionally, after a huge page is instantiated the PagePrivate flag is
+cleared so that accounting when the page is ultimately freed is correct.
+
+However, there are several instances where errors are encountered after a huge
+page is allocated but before it is instantiated. In this case, the page
+allocation has consumed the reservation and made the appropriate subpool,
+reservation map and global count adjustments. If the page is freed at this
+time (before instantiation and clearing of PagePrivate), then free_huge_page
+will increment the global reservation count. However, the reservation map
+indicates the reservation was consumed. This resulting inconsistent state
+will cause the 'leak' of a reserved huge page. The global reserve count will
+be higher than it should and prevent allocation of a pre-allocated page.
+
+The routine restore_reserve_on_error() attempts to handle this situation. It
+is fairly well documented. The intention of this routine is to restore
+the reservation map to the way it was before the page allocation. In this
+way, the state of the reservation map will correspond to the global reservation
+count after the page is freed.
+
+The routine restore_reserve_on_error itself may encounter errors while
+attempting to restore the reservation map entry. In this case, it will
+simply clear the PagePrivate flag of the page. In this way, the global
+reserve count will not be incremented when the page is freed. However, the
+reservation map will continue to look as though the reservation was consumed.
+A page can still be allocated for the address, but it will not use a reserved
+page as originally intended.
+
+There is some code (most notably userfaultfd) which can not call
+restore_reserve_on_error. In this case, it simply modifies the PagePrivate
+so that a reservation will not be leaked when the huge page is freed.
+
+
+Reservations and Memory Policy
+==============================
+Per-node huge page lists existed in struct hstate when git was first used
+to manage Linux code. The concept of reservations was added some time later.
+When reservations were added, no attempt was made to take memory policy
+into account. While cpusets are not exactly the same as memory policy, this
+comment in hugetlb_acct_memory sums up the interaction between reservations
+and cpusets/memory policy::
+
+ /*
+ * When cpuset is configured, it breaks the strict hugetlb page
+ * reservation as the accounting is done on a global variable. Such
+ * reservation is completely rubbish in the presence of cpuset because
+ * the reservation is not checked against page availability for the
+ * current cpuset. Application can still potentially OOM'ed by kernel
+ * with lack of free htlb page in cpuset that the task is in.
+ * Attempt to enforce strict accounting with cpuset is almost
+ * impossible (or too ugly) because cpuset is too fluid that
+ * task or memory node can be dynamically moved between cpusets.
+ *
+ * The change of semantics for shared hugetlb mapping with cpuset is
+ * undesirable. However, in order to preserve some of the semantics,
+ * we fall back to check against current free page availability as
+ * a best attempt and hopefully to minimize the impact of changing
+ * semantics that cpuset has.
+ */
+
+Huge page reservations were added to prevent unexpected page allocation
+failures (OOM) at page fault time. However, if an application makes use
+of cpusets or memory policy there is no guarantee that huge pages will be
+available on the required nodes. This is true even if there are a sufficient
+number of global reservations.
+
+Hugetlbfs regression testing
+============================
+
+The most complete set of hugetlb tests are in the libhugetlbfs repository.
+If you modify any hugetlb related code, use the libhugetlbfs test suite
+to check for regressions. In addition, if you add any new hugetlb
+functionality, please add appropriate tests to libhugetlbfs.
+
+--
+Mike Kravetz, 7 April 2017
diff --git a/Documentation/vm/hwpoison.rst b/Documentation/vm/hwpoison.rst
new file mode 100644
index 000000000..a5c884293
--- /dev/null
+++ b/Documentation/vm/hwpoison.rst
@@ -0,0 +1,186 @@
+.. hwpoison:
+
+========
+hwpoison
+========
+
+What is hwpoison?
+=================
+
+Upcoming Intel CPUs have support for recovering from some memory errors
+(``MCA recovery``). This requires the OS to declare a page "poisoned",
+kill the processes associated with it and avoid using it in the future.
+
+This patchkit implements the necessary infrastructure in the VM.
+
+To quote the overview comment::
+
+ High level machine check handler. Handles pages reported by the
+ hardware as being corrupted usually due to a 2bit ECC memory or cache
+ failure.
+
+ This focusses on pages detected as corrupted in the background.
+ When the current CPU tries to consume corruption the currently
+ running process can just be killed directly instead. This implies
+ that if the error cannot be handled for some reason it's safe to
+ just ignore it because no corruption has been consumed yet. Instead
+ when that happens another machine check will happen.
+
+ Handles page cache pages in various states. The tricky part
+ here is that we can access any page asynchronous to other VM
+ users, because memory failures could happen anytime and anywhere,
+ possibly violating some of their assumptions. This is why this code
+ has to be extremely careful. Generally it tries to use normal locking
+ rules, as in get the standard locks, even if that means the
+ error handling takes potentially a long time.
+
+ Some of the operations here are somewhat inefficient and have non
+ linear algorithmic complexity, because the data structures have not
+ been optimized for this case. This is in particular the case
+ for the mapping from a vma to a process. Since this case is expected
+ to be rare we hope we can get away with this.
+
+The code consists of a the high level handler in mm/memory-failure.c,
+a new page poison bit and various checks in the VM to handle poisoned
+pages.
+
+The main target right now is KVM guests, but it works for all kinds
+of applications. KVM support requires a recent qemu-kvm release.
+
+For the KVM use there was need for a new signal type so that
+KVM can inject the machine check into the guest with the proper
+address. This in theory allows other applications to handle
+memory failures too. The expection is that near all applications
+won't do that, but some very specialized ones might.
+
+Failure recovery modes
+======================
+
+There are two (actually three) modes memory failure recovery can be in:
+
+vm.memory_failure_recovery sysctl set to zero:
+ All memory failures cause a panic. Do not attempt recovery.
+ (on x86 this can be also affected by the tolerant level of the
+ MCE subsystem)
+
+early kill
+ (can be controlled globally and per process)
+ Send SIGBUS to the application as soon as the error is detected
+ This allows applications who can process memory errors in a gentle
+ way (e.g. drop affected object)
+ This is the mode used by KVM qemu.
+
+late kill
+ Send SIGBUS when the application runs into the corrupted page.
+ This is best for memory error unaware applications and default
+ Note some pages are always handled as late kill.
+
+User control
+============
+
+vm.memory_failure_recovery
+ See sysctl.txt
+
+vm.memory_failure_early_kill
+ Enable early kill mode globally
+
+PR_MCE_KILL
+ Set early/late kill mode/revert to system default
+
+ arg1: PR_MCE_KILL_CLEAR:
+ Revert to system default
+ arg1: PR_MCE_KILL_SET:
+ arg2 defines thread specific mode
+
+ PR_MCE_KILL_EARLY:
+ Early kill
+ PR_MCE_KILL_LATE:
+ Late kill
+ PR_MCE_KILL_DEFAULT
+ Use system global default
+
+ Note that if you want to have a dedicated thread which handles
+ the SIGBUS(BUS_MCEERR_AO) on behalf of the process, you should
+ call prctl(PR_MCE_KILL_EARLY) on the designated thread. Otherwise,
+ the SIGBUS is sent to the main thread.
+
+PR_MCE_KILL_GET
+ return current mode
+
+Testing
+=======
+
+* madvise(MADV_HWPOISON, ....) (as root) - Poison a page in the
+ process for testing
+
+* hwpoison-inject module through debugfs ``/sys/kernel/debug/hwpoison/``
+
+ corrupt-pfn
+ Inject hwpoison fault at PFN echoed into this file. This does
+ some early filtering to avoid corrupted unintended pages in test suites.
+
+ unpoison-pfn
+ Software-unpoison page at PFN echoed into this file. This way
+ a page can be reused again. This only works for Linux
+ injected failures, not for real memory failures.
+
+ Note these injection interfaces are not stable and might change between
+ kernel versions
+
+ corrupt-filter-dev-major, corrupt-filter-dev-minor
+ Only handle memory failures to pages associated with the file
+ system defined by block device major/minor. -1U is the
+ wildcard value. This should be only used for testing with
+ artificial injection.
+
+ corrupt-filter-memcg
+ Limit injection to pages owned by memgroup. Specified by inode
+ number of the memcg.
+
+ Example::
+
+ mkdir /sys/fs/cgroup/mem/hwpoison
+
+ usemem -m 100 -s 1000 &
+ echo `jobs -p` > /sys/fs/cgroup/mem/hwpoison/tasks
+
+ memcg_ino=$(ls -id /sys/fs/cgroup/mem/hwpoison | cut -f1 -d' ')
+ echo $memcg_ino > /debug/hwpoison/corrupt-filter-memcg
+
+ page-types -p `pidof init` --hwpoison # shall do nothing
+ page-types -p `pidof usemem` --hwpoison # poison its pages
+
+ corrupt-filter-flags-mask, corrupt-filter-flags-value
+ When specified, only poison pages if ((page_flags & mask) ==
+ value). This allows stress testing of many kinds of
+ pages. The page_flags are the same as in /proc/kpageflags. The
+ flag bits are defined in include/linux/kernel-page-flags.h and
+ documented in Documentation/admin-guide/mm/pagemap.rst
+
+* Architecture specific MCE injector
+
+ x86 has mce-inject, mce-test
+
+ Some portable hwpoison test programs in mce-test, see below.
+
+References
+==========
+
+http://halobates.de/mce-lc09-2.pdf
+ Overview presentation from LinuxCon 09
+
+git://git.kernel.org/pub/scm/utils/cpu/mce/mce-test.git
+ Test suite (hwpoison specific portable tests in tsrc)
+
+git://git.kernel.org/pub/scm/utils/cpu/mce/mce-inject.git
+ x86 specific injector
+
+
+Limitations
+===========
+- Not all page types are supported and never will. Most kernel internal
+ objects cannot be recovered, only LRU pages for now.
+- Right now hugepage support is missing.
+
+---
+Andi Kleen, Oct 2009
diff --git a/Documentation/vm/index.rst b/Documentation/vm/index.rst
new file mode 100644
index 000000000..eff5fbd49
--- /dev/null
+++ b/Documentation/vm/index.rst
@@ -0,0 +1,55 @@
+=====================================
+Linux Memory Management Documentation
+=====================================
+
+This is a collection of documents about the Linux memory management (mm)
+subsystem. If you are looking for advice on simply allocating memory,
+see the :ref:`memory_allocation`.
+
+User guides for MM features
+===========================
+
+The following documents provide guides for controlling and tuning
+various features of the Linux memory management
+
+.. toctree::
+ :maxdepth: 1
+
+ swap_numa
+ zswap
+
+Kernel developers MM documentation
+==================================
+
+The below documents describe MM internals with different level of
+details ranging from notes and mailing list responses to elaborate
+descriptions of data structures and algorithms.
+
+.. toctree::
+ :maxdepth: 1
+
+ active_mm
+ arch_pgtable_helpers
+ balance
+ cleancache
+ free_page_reporting
+ frontswap
+ highmem
+ hmm
+ hwpoison
+ hugetlbfs_reserv
+ ksm
+ memory-model
+ mmu_notifier
+ numa
+ overcommit-accounting
+ page_migration
+ page_frags
+ page_owner
+ remap_file_pages
+ slub
+ split_page_table_lock
+ transhuge
+ unevictable-lru
+ z3fold
+ zsmalloc
diff --git a/Documentation/vm/ksm.rst b/Documentation/vm/ksm.rst
new file mode 100644
index 000000000..9e37add06
--- /dev/null
+++ b/Documentation/vm/ksm.rst
@@ -0,0 +1,87 @@
+.. _ksm:
+
+=======================
+Kernel Samepage Merging
+=======================
+
+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 https://lwn.net/Articles/330589/
+
+The userspace interface of KSM is described in :ref:`Documentation/admin-guide/mm/ksm.rst <admin_guide_ksm>`
+
+Design
+======
+
+Overview
+--------
+
+.. kernel-doc:: mm/ksm.c
+ :DOC: Overview
+
+Reverse mapping
+---------------
+KSM maintains reverse mapping information for KSM pages in the stable
+tree.
+
+If a KSM page is shared between less than ``max_page_sharing`` VMAs,
+the node of the stable tree that represents such KSM page points to a
+list of struct rmap_item and the ``page->mapping`` of the
+KSM page points to the stable tree node.
+
+When the sharing passes this threshold, KSM adds a second dimension to
+the stable tree. The tree node becomes a "chain" that links one or
+more "dups". Each "dup" keeps reverse mapping information for a KSM
+page with ``page->mapping`` pointing to that "dup".
+
+Every "chain" and all "dups" linked into a "chain" enforce the
+invariant that they represent the same write protected memory content,
+even if each "dup" will be pointed by a different KSM page copy of
+that content.
+
+This way the stable tree lookup computational complexity is unaffected
+if compared to an unlimited list of reverse mappings. It is still
+enforced that there cannot be KSM page content duplicates in the
+stable tree itself.
+
+The deduplication limit enforced by ``max_page_sharing`` is required
+to avoid the virtual memory rmap lists to grow too large. The rmap
+walk has O(N) complexity where N is the number of rmap_items
+(i.e. virtual mappings) that are sharing the page, which is in turn
+capped by ``max_page_sharing``. So this effectively spreads the linear
+O(N) computational complexity from rmap walk context over different
+KSM pages. The ksmd walk over the stable_node "chains" is also O(N),
+but N is the number of stable_node "dups", not the number of
+rmap_items, so it has not a significant impact on ksmd performance. In
+practice the best stable_node "dup" candidate will be kept and found
+at the head of the "dups" list.
+
+High values of ``max_page_sharing`` result in faster memory merging
+(because there will be fewer stable_node dups queued into the
+stable_node chain->hlist to check for pruning) and higher
+deduplication factor at the expense of slower worst case for rmap
+walks for any KSM page which can happen during swapping, compaction,
+NUMA balancing and page migration.
+
+The ``stable_node_dups/stable_node_chains`` ratio is also affected by the
+``max_page_sharing`` tunable, and an high ratio may indicate fragmentation
+in the stable_node dups, which could be solved by introducing
+fragmentation algorithms in ksmd which would refile rmap_items from
+one stable_node dup to another stable_node dup, in order to free up
+stable_node "dups" with few rmap_items in them, but that may increase
+the ksmd CPU usage and possibly slowdown the readonly computations on
+the KSM pages of the applications.
+
+The whole list of stable_node "dups" linked in the stable_node
+"chains" is scanned periodically in order to prune stale stable_nodes.
+The frequency of such scans is defined by
+``stable_node_chains_prune_millisecs`` sysfs tunable.
+
+Reference
+---------
+.. kernel-doc:: mm/ksm.c
+ :functions: mm_slot ksm_scan stable_node rmap_item
+
+--
+Izik Eidus,
+Hugh Dickins, 17 Nov 2009
diff --git a/Documentation/vm/memory-model.rst b/Documentation/vm/memory-model.rst
new file mode 100644
index 000000000..ce398a7dc
--- /dev/null
+++ b/Documentation/vm/memory-model.rst
@@ -0,0 +1,218 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+.. _physical_memory_model:
+
+=====================
+Physical Memory Model
+=====================
+
+Physical memory in a system may be addressed in different ways. The
+simplest case is when the physical memory starts at address 0 and
+spans a contiguous range up to the maximal address. It could be,
+however, that this range contains small holes that are not accessible
+for the CPU. Then there could be several contiguous ranges at
+completely distinct addresses. And, don't forget about NUMA, where
+different memory banks are attached to different CPUs.
+
+Linux abstracts this diversity using one of the three memory models:
+FLATMEM, DISCONTIGMEM and SPARSEMEM. Each architecture defines what
+memory models it supports, what the default memory model is and
+whether it is possible to manually override that default.
+
+.. note::
+ At time of this writing, DISCONTIGMEM is considered deprecated,
+ although it is still in use by several architectures.
+
+All the memory models track the status of physical page frames using
+struct page arranged in one or more arrays.
+
+Regardless of the selected memory model, there exists one-to-one
+mapping between the physical page frame number (PFN) and the
+corresponding `struct page`.
+
+Each memory model defines :c:func:`pfn_to_page` and :c:func:`page_to_pfn`
+helpers that allow the conversion from PFN to `struct page` and vice
+versa.
+
+FLATMEM
+=======
+
+The simplest memory model is FLATMEM. This model is suitable for
+non-NUMA systems with contiguous, or mostly contiguous, physical
+memory.
+
+In the FLATMEM memory model, there is a global `mem_map` array that
+maps the entire physical memory. For most architectures, the holes
+have entries in the `mem_map` array. The `struct page` objects
+corresponding to the holes are never fully initialized.
+
+To allocate the `mem_map` array, architecture specific setup code should
+call :c:func:`free_area_init` function. Yet, the mappings array is not
+usable until the call to :c:func:`memblock_free_all` that hands all the
+memory to the page allocator.
+
+An architecture may free parts of the `mem_map` array that do not cover the
+actual physical pages. In such case, the architecture specific
+:c:func:`pfn_valid` implementation should take the holes in the
+`mem_map` into account.
+
+With FLATMEM, the conversion between a PFN and the `struct page` is
+straightforward: `PFN - ARCH_PFN_OFFSET` is an index to the
+`mem_map` array.
+
+The `ARCH_PFN_OFFSET` defines the first page frame number for
+systems with physical memory starting at address different from 0.
+
+DISCONTIGMEM
+============
+
+The DISCONTIGMEM model treats the physical memory as a collection of
+`nodes` similarly to how Linux NUMA support does. For each node Linux
+constructs an independent memory management subsystem represented by
+`struct pglist_data` (or `pg_data_t` for short). Among other
+things, `pg_data_t` holds the `node_mem_map` array that maps
+physical pages belonging to that node. The `node_start_pfn` field of
+`pg_data_t` is the number of the first page frame belonging to that
+node.
+
+The architecture setup code should call :c:func:`free_area_init_node` for
+each node in the system to initialize the `pg_data_t` object and its
+`node_mem_map`.
+
+Every `node_mem_map` behaves exactly as FLATMEM's `mem_map` -
+every physical page frame in a node has a `struct page` entry in the
+`node_mem_map` array. When DISCONTIGMEM is enabled, a portion of the
+`flags` field of the `struct page` encodes the node number of the
+node hosting that page.
+
+The conversion between a PFN and the `struct page` in the
+DISCONTIGMEM model became slightly more complex as it has to determine
+which node hosts the physical page and which `pg_data_t` object
+holds the `struct page`.
+
+Architectures that support DISCONTIGMEM provide :c:func:`pfn_to_nid`
+to convert PFN to the node number. The opposite conversion helper
+:c:func:`page_to_nid` is generic as it uses the node number encoded in
+page->flags.
+
+Once the node number is known, the PFN can be used to index
+appropriate `node_mem_map` array to access the `struct page` and
+the offset of the `struct page` from the `node_mem_map` plus
+`node_start_pfn` is the PFN of that page.
+
+SPARSEMEM
+=========
+
+SPARSEMEM is the most versatile memory model available in Linux and it
+is the only memory model that supports several advanced features such
+as hot-plug and hot-remove of the physical memory, alternative memory
+maps for non-volatile memory devices and deferred initialization of
+the memory map for larger systems.
+
+The SPARSEMEM model presents the physical memory as a collection of
+sections. A section is represented with struct mem_section
+that contains `section_mem_map` that is, logically, a pointer to an
+array of struct pages. However, it is stored with some other magic
+that aids the sections management. The section size and maximal number
+of section is specified using `SECTION_SIZE_BITS` and
+`MAX_PHYSMEM_BITS` constants defined by each architecture that
+supports SPARSEMEM. While `MAX_PHYSMEM_BITS` is an actual width of a
+physical address that an architecture supports, the
+`SECTION_SIZE_BITS` is an arbitrary value.
+
+The maximal number of sections is denoted `NR_MEM_SECTIONS` and
+defined as
+
+.. math::
+
+ NR\_MEM\_SECTIONS = 2 ^ {(MAX\_PHYSMEM\_BITS - SECTION\_SIZE\_BITS)}
+
+The `mem_section` objects are arranged in a two-dimensional array
+called `mem_sections`. The size and placement of this array depend
+on `CONFIG_SPARSEMEM_EXTREME` and the maximal possible number of
+sections:
+
+* When `CONFIG_SPARSEMEM_EXTREME` is disabled, the `mem_sections`
+ array is static and has `NR_MEM_SECTIONS` rows. Each row holds a
+ single `mem_section` object.
+* When `CONFIG_SPARSEMEM_EXTREME` is enabled, the `mem_sections`
+ array is dynamically allocated. Each row contains PAGE_SIZE worth of
+ `mem_section` objects and the number of rows is calculated to fit
+ all the memory sections.
+
+The architecture setup code should call sparse_init() to
+initialize the memory sections and the memory maps.
+
+With SPARSEMEM there are two possible ways to convert a PFN to the
+corresponding `struct page` - a "classic sparse" and "sparse
+vmemmap". The selection is made at build time and it is determined by
+the value of `CONFIG_SPARSEMEM_VMEMMAP`.
+
+The classic sparse encodes the section number of a page in page->flags
+and uses high bits of a PFN to access the section that maps that page
+frame. Inside a section, the PFN is the index to the array of pages.
+
+The sparse vmemmap uses a virtually mapped memory map to optimize
+pfn_to_page and page_to_pfn operations. There is a global `struct
+page *vmemmap` pointer that points to a virtually contiguous array of
+`struct page` objects. A PFN is an index to that array and the
+offset of the `struct page` from `vmemmap` is the PFN of that
+page.
+
+To use vmemmap, an architecture has to reserve a range of virtual
+addresses that will map the physical pages containing the memory
+map and make sure that `vmemmap` points to that range. In addition,
+the architecture should implement :c:func:`vmemmap_populate` method
+that will allocate the physical memory and create page tables for the
+virtual memory map. If an architecture does not have any special
+requirements for the vmemmap mappings, it can use default
+:c:func:`vmemmap_populate_basepages` provided by the generic memory
+management.
+
+The virtually mapped memory map allows storing `struct page` objects
+for persistent memory devices in pre-allocated storage on those
+devices. This storage is represented with struct vmem_altmap
+that is eventually passed to vmemmap_populate() through a long chain
+of function calls. The vmemmap_populate() implementation may use the
+`vmem_altmap` along with :c:func:`vmemmap_alloc_block_buf` helper to
+allocate memory map on the persistent memory device.
+
+ZONE_DEVICE
+===========
+The `ZONE_DEVICE` facility builds upon `SPARSEMEM_VMEMMAP` to offer
+`struct page` `mem_map` services for device driver identified physical
+address ranges. The "device" aspect of `ZONE_DEVICE` relates to the fact
+that the page objects for these address ranges are never marked online,
+and that a reference must be taken against the device, not just the page
+to keep the memory pinned for active use. `ZONE_DEVICE`, via
+:c:func:`devm_memremap_pages`, performs just enough memory hotplug to
+turn on :c:func:`pfn_to_page`, :c:func:`page_to_pfn`, and
+:c:func:`get_user_pages` service for the given range of pfns. Since the
+page reference count never drops below 1 the page is never tracked as
+free memory and the page's `struct list_head lru` space is repurposed
+for back referencing to the host device / driver that mapped the memory.
+
+While `SPARSEMEM` presents memory as a collection of sections,
+optionally collected into memory blocks, `ZONE_DEVICE` users have a need
+for smaller granularity of populating the `mem_map`. Given that
+`ZONE_DEVICE` memory is never marked online it is subsequently never
+subject to its memory ranges being exposed through the sysfs memory
+hotplug api on memory block boundaries. The implementation relies on
+this lack of user-api constraint to allow sub-section sized memory
+ranges to be specified to :c:func:`arch_add_memory`, the top-half of
+memory hotplug. Sub-section support allows for 2MB as the cross-arch
+common alignment granularity for :c:func:`devm_memremap_pages`.
+
+The users of `ZONE_DEVICE` are:
+
+* pmem: Map platform persistent memory to be used as a direct-I/O target
+ via DAX mappings.
+
+* hmm: Extend `ZONE_DEVICE` with `->page_fault()` and `->page_free()`
+ event callbacks to allow a device-driver to coordinate memory management
+ events related to device-memory, typically GPU memory. See
+ Documentation/vm/hmm.rst.
+
+* p2pdma: Create `struct page` objects to allow peer devices in a
+ PCI/-E topology to coordinate direct-DMA operations between themselves,
+ i.e. bypass host memory.
diff --git a/Documentation/vm/mmu_notifier.rst b/Documentation/vm/mmu_notifier.rst
new file mode 100644
index 000000000..df5d7777f
--- /dev/null
+++ b/Documentation/vm/mmu_notifier.rst
@@ -0,0 +1,99 @@
+.. _mmu_notifier:
+
+When do you need to notify inside page table lock ?
+===================================================
+
+When clearing a pte/pmd we are given a choice to notify the event through
+(notify version of \*_clear_flush call mmu_notifier_invalidate_range) under
+the page table lock. But that notification is not necessary in all cases.
+
+For secondary TLB (non CPU TLB) like IOMMU TLB or device TLB (when device use
+thing like ATS/PASID to get the IOMMU to walk the CPU page table to access a
+process virtual address space). There is only 2 cases when you need to notify
+those secondary TLB while holding page table lock when clearing a pte/pmd:
+
+ A) page backing address is free before mmu_notifier_invalidate_range_end()
+ B) a page table entry is updated to point to a new page (COW, write fault
+ on zero page, __replace_page(), ...)
+
+Case A is obvious you do not want to take the risk for the device to write to
+a page that might now be used by some completely different task.
+
+Case B is more subtle. For correctness it requires the following sequence to
+happen:
+
+ - take page table lock
+ - clear page table entry and notify ([pmd/pte]p_huge_clear_flush_notify())
+ - set page table entry to point to new page
+
+If clearing the page table entry is not followed by a notify before setting
+the new pte/pmd value then you can break memory model like C11 or C++11 for
+the device.
+
+Consider the following scenario (device use a feature similar to ATS/PASID):
+
+Two address addrA and addrB such that \|addrA - addrB\| >= PAGE_SIZE we assume
+they are write protected for COW (other case of B apply too).
+
+::
+
+ [Time N] --------------------------------------------------------------------
+ CPU-thread-0 {try to write to addrA}
+ CPU-thread-1 {try to write to addrB}
+ CPU-thread-2 {}
+ CPU-thread-3 {}
+ DEV-thread-0 {read addrA and populate device TLB}
+ DEV-thread-2 {read addrB and populate device TLB}
+ [Time N+1] ------------------------------------------------------------------
+ CPU-thread-0 {COW_step0: {mmu_notifier_invalidate_range_start(addrA)}}
+ CPU-thread-1 {COW_step0: {mmu_notifier_invalidate_range_start(addrB)}}
+ CPU-thread-2 {}
+ CPU-thread-3 {}
+ DEV-thread-0 {}
+ DEV-thread-2 {}
+ [Time N+2] ------------------------------------------------------------------
+ CPU-thread-0 {COW_step1: {update page table to point to new page for addrA}}
+ CPU-thread-1 {COW_step1: {update page table to point to new page for addrB}}
+ CPU-thread-2 {}
+ CPU-thread-3 {}
+ DEV-thread-0 {}
+ DEV-thread-2 {}
+ [Time N+3] ------------------------------------------------------------------
+ CPU-thread-0 {preempted}
+ CPU-thread-1 {preempted}
+ CPU-thread-2 {write to addrA which is a write to new page}
+ CPU-thread-3 {}
+ DEV-thread-0 {}
+ DEV-thread-2 {}
+ [Time N+3] ------------------------------------------------------------------
+ CPU-thread-0 {preempted}
+ CPU-thread-1 {preempted}
+ CPU-thread-2 {}
+ CPU-thread-3 {write to addrB which is a write to new page}
+ DEV-thread-0 {}
+ DEV-thread-2 {}
+ [Time N+4] ------------------------------------------------------------------
+ CPU-thread-0 {preempted}
+ CPU-thread-1 {COW_step3: {mmu_notifier_invalidate_range_end(addrB)}}
+ CPU-thread-2 {}
+ CPU-thread-3 {}
+ DEV-thread-0 {}
+ DEV-thread-2 {}
+ [Time N+5] ------------------------------------------------------------------
+ CPU-thread-0 {preempted}
+ CPU-thread-1 {}
+ CPU-thread-2 {}
+ CPU-thread-3 {}
+ DEV-thread-0 {read addrA from old page}
+ DEV-thread-2 {read addrB from new page}
+
+So here because at time N+2 the clear page table entry was not pair with a
+notification to invalidate the secondary TLB, the device see the new value for
+addrB before seeing the new value for addrA. This break total memory ordering
+for the device.
+
+When changing a pte to write protect or to point to a new write protected page
+with same content (KSM) it is fine to delay the mmu_notifier_invalidate_range
+call to mmu_notifier_invalidate_range_end() outside the page table lock. This
+is true even if the thread doing the page table update is preempted right after
+releasing page table lock but before call mmu_notifier_invalidate_range_end().
diff --git a/Documentation/vm/numa.rst b/Documentation/vm/numa.rst
new file mode 100644
index 000000000..99fdeca91
--- /dev/null
+++ b/Documentation/vm/numa.rst
@@ -0,0 +1,150 @@
+.. _numa:
+
+Started Nov 1999 by Kanoj Sarcar <kanoj@sgi.com>
+
+=============
+What is NUMA?
+=============
+
+This question can be answered from a couple of perspectives: the
+hardware view and the Linux software view.
+
+From the hardware perspective, a NUMA system is a computer platform that
+comprises multiple components or assemblies each of which may contain 0
+or more CPUs, local memory, and/or IO buses. For brevity and to
+disambiguate the hardware view of these physical components/assemblies
+from the software abstraction thereof, we'll call the components/assemblies
+'cells' in this document.
+
+Each of the 'cells' may be viewed as an SMP [symmetric multi-processor] subset
+of the system--although some components necessary for a stand-alone SMP system
+may not be populated on any given cell. The cells of the NUMA system are
+connected together with some sort of system interconnect--e.g., a crossbar or
+point-to-point link are common types of NUMA system interconnects. Both of
+these types of interconnects can be aggregated to create NUMA platforms with
+cells at multiple distances from other cells.
+
+For Linux, the NUMA platforms of interest are primarily what is known as Cache
+Coherent NUMA or ccNUMA systems. With ccNUMA systems, all memory is visible
+to and accessible from any CPU attached to any cell and cache coherency
+is handled in hardware by the processor caches and/or the system interconnect.
+
+Memory access time and effective memory bandwidth varies depending on how far
+away the cell containing the CPU or IO bus making the memory access is from the
+cell containing the target memory. For example, access to memory by CPUs
+attached to the same cell will experience faster access times and higher
+bandwidths than accesses to memory on other, remote cells. NUMA platforms
+can have cells at multiple remote distances from any given cell.
+
+Platform vendors don't build NUMA systems just to make software developers'
+lives interesting. Rather, this architecture is a means to provide scalable
+memory bandwidth. However, to achieve scalable memory bandwidth, system and
+application software must arrange for a large majority of the memory references
+[cache misses] to be to "local" memory--memory on the same cell, if any--or
+to the closest cell with memory.
+
+This leads to the Linux software view of a NUMA system:
+
+Linux divides the system's hardware resources into multiple software
+abstractions called "nodes". Linux maps the nodes onto the physical cells
+of the hardware platform, abstracting away some of the details for some
+architectures. As with physical cells, software nodes may contain 0 or more
+CPUs, memory and/or IO buses. And, again, memory accesses to memory on
+"closer" nodes--nodes that map to closer cells--will generally experience
+faster access times and higher effective bandwidth than accesses to more
+remote cells.
+
+For some architectures, such as x86, Linux will "hide" any node representing a
+physical cell that has no memory attached, and reassign any CPUs attached to
+that cell to a node representing a cell that does have memory. Thus, on
+these architectures, one cannot assume that all CPUs that Linux associates with
+a given node will see the same local memory access times and bandwidth.
+
+In addition, for some architectures, again x86 is an example, Linux supports
+the emulation of additional nodes. For NUMA emulation, linux will carve up
+the existing nodes--or the system memory for non-NUMA platforms--into multiple
+nodes. Each emulated node will manage a fraction of the underlying cells'
+physical memory. NUMA emluation is useful for testing NUMA kernel and
+application features on non-NUMA platforms, and as a sort of memory resource
+management mechanism when used together with cpusets.
+[see Documentation/admin-guide/cgroup-v1/cpusets.rst]
+
+For each node with memory, Linux constructs an independent memory management
+subsystem, complete with its own free page lists, in-use page lists, usage
+statistics and locks to mediate access. In addition, Linux constructs for
+each memory zone [one or more of DMA, DMA32, NORMAL, HIGH_MEMORY, MOVABLE],
+an ordered "zonelist". A zonelist specifies the zones/nodes to visit when a
+selected zone/node cannot satisfy the allocation request. This situation,
+when a zone has no available memory to satisfy a request, is called
+"overflow" or "fallback".
+
+Because some nodes contain multiple zones containing different types of
+memory, Linux must decide whether to order the zonelists such that allocations
+fall back to the same zone type on a different node, or to a different zone
+type on the same node. This is an important consideration because some zones,
+such as DMA or DMA32, represent relatively scarce resources. Linux chooses
+a default Node ordered zonelist. This means it tries to fallback to other zones
+from the same node before using remote nodes which are ordered by NUMA distance.
+
+By default, Linux will attempt to satisfy memory allocation requests from the
+node to which the CPU that executes the request is assigned. Specifically,
+Linux will attempt to allocate from the first node in the appropriate zonelist
+for the node where the request originates. This is called "local allocation."
+If the "local" node cannot satisfy the request, the kernel will examine other
+nodes' zones in the selected zonelist looking for the first zone in the list
+that can satisfy the request.
+
+Local allocation will tend to keep subsequent access to the allocated memory
+"local" to the underlying physical resources and off the system interconnect--
+as long as the task on whose behalf the kernel allocated some memory does not
+later migrate away from that memory. The Linux scheduler is aware of the
+NUMA topology of the platform--embodied in the "scheduling domains" data
+structures [see Documentation/scheduler/sched-domains.rst]--and the scheduler
+attempts to minimize task migration to distant scheduling domains. However,
+the scheduler does not take a task's NUMA footprint into account directly.
+Thus, under sufficient imbalance, tasks can migrate between nodes, remote
+from their initial node and kernel data structures.
+
+System administrators and application designers can restrict a task's migration
+to improve NUMA locality using various CPU affinity command line interfaces,
+such as taskset(1) and numactl(1), and program interfaces such as
+sched_setaffinity(2). Further, one can modify the kernel's default local
+allocation behavior using Linux NUMA memory policy. [see
+:ref:`Documentation/admin-guide/mm/numa_memory_policy.rst <numa_memory_policy>`].
+
+System administrators can restrict the CPUs and nodes' memories that a non-
+privileged user can specify in the scheduling or NUMA commands and functions
+using control groups and CPUsets. [see Documentation/admin-guide/cgroup-v1/cpusets.rst]
+
+On architectures that do not hide memoryless nodes, Linux will include only
+zones [nodes] with memory in the zonelists. This means that for a memoryless
+node the "local memory node"--the node of the first zone in CPU's node's
+zonelist--will not be the node itself. Rather, it will be the node that the
+kernel selected as the nearest node with memory when it built the zonelists.
+So, default, local allocations will succeed with the kernel supplying the
+closest available memory. This is a consequence of the same mechanism that
+allows such allocations to fallback to other nearby nodes when a node that
+does contain memory overflows.
+
+Some kernel allocations do not want or cannot tolerate this allocation fallback
+behavior. Rather they want to be sure they get memory from the specified node
+or get notified that the node has no free memory. This is usually the case when
+a subsystem allocates per CPU memory resources, for example.
+
+A typical model for making such an allocation is to obtain the node id of the
+node to which the "current CPU" is attached using one of the kernel's
+numa_node_id() or CPU_to_node() functions and then request memory from only
+the node id returned. When such an allocation fails, the requesting subsystem
+may revert to its own fallback path. The slab kernel memory allocator is an
+example of this. Or, the subsystem may choose to disable or not to enable
+itself on allocation failure. The kernel profiling subsystem is an example of
+this.
+
+If the architecture supports--does not hide--memoryless nodes, then CPUs
+attached to memoryless nodes would always incur the fallback path overhead
+or some subsystems would fail to initialize if they attempted to allocated
+memory exclusively from a node without memory. To support such
+architectures transparently, kernel subsystems can use the numa_mem_id()
+or cpu_to_mem() function to locate the "local memory node" for the calling or
+specified CPU. Again, this is the same node from which default, local page
+allocations will be attempted.
diff --git a/Documentation/vm/overcommit-accounting.rst b/Documentation/vm/overcommit-accounting.rst
new file mode 100644
index 000000000..0dd54bbe4
--- /dev/null
+++ b/Documentation/vm/overcommit-accounting.rst
@@ -0,0 +1,87 @@
+.. _overcommit_accounting:
+
+=====================
+Overcommit Accounting
+=====================
+
+The Linux kernel supports the following overcommit handling modes
+
+0
+ Heuristic overcommit handling. Obvious overcommits of address
+ space are refused. Used for a typical system. It ensures a
+ seriously wild allocation fails while allowing overcommit to
+ reduce swap usage. root is allowed to allocate slightly more
+ memory in this mode. This is the default.
+
+1
+ Always overcommit. Appropriate for some scientific
+ applications. Classic example is code using sparse arrays and
+ just relying on the virtual memory consisting almost entirely
+ of zero pages.
+
+2
+ Don't overcommit. The total address space commit for the
+ system is not permitted to exceed swap + a configurable amount
+ (default is 50%) of physical RAM. Depending on the amount you
+ use, in most situations this means a process will not be
+ killed while accessing pages but will receive errors on memory
+ allocation as appropriate.
+
+ Useful for applications that want to guarantee their memory
+ allocations will be available in the future without having to
+ initialize every page.
+
+The overcommit policy is set via the sysctl ``vm.overcommit_memory``.
+
+The overcommit amount can be set via ``vm.overcommit_ratio`` (percentage)
+or ``vm.overcommit_kbytes`` (absolute value).
+
+The current overcommit limit and amount committed are viewable in
+``/proc/meminfo`` as CommitLimit and Committed_AS respectively.
+
+Gotchas
+=======
+
+The C language stack growth does an implicit mremap. If you want absolute
+guarantees and run close to the edge you MUST mmap your stack for the
+largest size you think you will need. For typical stack usage this does
+not matter much but it's a corner case if you really really care
+
+In mode 2 the MAP_NORESERVE flag is ignored.
+
+
+How It Works
+============
+
+The overcommit is based on the following rules
+
+For a file backed map
+ | SHARED or READ-only - 0 cost (the file is the map not swap)
+ | PRIVATE WRITABLE - size of mapping per instance
+
+For an anonymous or ``/dev/zero`` map
+ | SHARED - size of mapping
+ | PRIVATE READ-only - 0 cost (but of little use)
+ | PRIVATE WRITABLE - size of mapping per instance
+
+Additional accounting
+ | Pages made writable copies by mmap
+ | shmfs memory drawn from the same pool
+
+Status
+======
+
+* We account mmap memory mappings
+* We account mprotect changes in commit
+* We account mremap changes in size
+* We account brk
+* We account munmap
+* We report the commit status in /proc
+* Account and check on fork
+* Review stack handling/building on exec
+* SHMfs accounting
+* Implement actual limit enforcement
+
+To Do
+=====
+* Account ptrace pages (this is hard)
diff --git a/Documentation/vm/page_frags.rst b/Documentation/vm/page_frags.rst
new file mode 100644
index 000000000..7d6f9385d
--- /dev/null
+++ b/Documentation/vm/page_frags.rst
@@ -0,0 +1,45 @@
+.. _page_frags:
+
+==============
+Page fragments
+==============
+
+A page fragment is an arbitrary-length arbitrary-offset area of memory
+which resides within a 0 or higher order compound page. Multiple
+fragments within that page are individually refcounted, in the page's
+reference counter.
+
+The page_frag functions, page_frag_alloc and page_frag_free, provide a
+simple allocation framework for page fragments. This is used by the
+network stack and network device drivers to provide a backing region of
+memory for use as either an sk_buff->head, or to be used in the "frags"
+portion of skb_shared_info.
+
+In order to make use of the page fragment APIs a backing page fragment
+cache is needed. This provides a central point for the fragment allocation
+and tracks allows multiple calls to make use of a cached page. The
+advantage to doing this is that multiple calls to get_page can be avoided
+which can be expensive at allocation time. However due to the nature of
+this caching it is required that any calls to the cache be protected by
+either a per-cpu limitation, or a per-cpu limitation and forcing interrupts
+to be disabled when executing the fragment allocation.
+
+The network stack uses two separate caches per CPU to handle fragment
+allocation. The netdev_alloc_cache is used by callers making use of the
+netdev_alloc_frag and __netdev_alloc_skb calls. The napi_alloc_cache is
+used by callers of the __napi_alloc_frag and __napi_alloc_skb calls. The
+main difference between these two calls is the context in which they may be
+called. The "netdev" prefixed functions are usable in any context as these
+functions will disable interrupts, while the "napi" prefixed functions are
+only usable within the softirq context.
+
+Many network device drivers use a similar methodology for allocating page
+fragments, but the page fragments are cached at the ring or descriptor
+level. In order to enable these cases it is necessary to provide a generic
+way of tearing down a page cache. For this reason __page_frag_cache_drain
+was implemented. It allows for freeing multiple references from a single
+page via a single call. The advantage to doing this is that it allows for
+cleaning up the multiple references that were added to a page in order to
+avoid calling get_page per allocation.
+
+Alexander Duyck, Nov 29, 2016.
diff --git a/Documentation/vm/page_migration.rst b/Documentation/vm/page_migration.rst
new file mode 100644
index 000000000..db9d7e553
--- /dev/null
+++ b/Documentation/vm/page_migration.rst
@@ -0,0 +1,288 @@
+.. _page_migration:
+
+==============
+Page migration
+==============
+
+Page migration allows moving the physical location of pages between
+nodes in a NUMA system while the process is running. This means that the
+virtual addresses that the process sees do not change. However, the
+system rearranges the physical location of those pages.
+
+Also see :ref:`Heterogeneous Memory Management (HMM) <hmm>`
+for migrating pages to or from device private memory.
+
+The main intent of page migration is to reduce the latency of memory accesses
+by moving pages near to the processor where the process accessing that memory
+is running.
+
+Page migration allows a process to manually relocate the node on which its
+pages are located through the MF_MOVE and MF_MOVE_ALL options while setting
+a new memory policy via mbind(). The pages of a process can also be relocated
+from another process using the sys_migrate_pages() function call. The
+migrate_pages() function call takes two sets of nodes and moves pages of a
+process that are located on the from nodes to the destination nodes.
+Page migration functions are provided by the numactl package by Andi Kleen
+(a version later than 0.9.3 is required. Get it from
+https://github.com/numactl/numactl.git). numactl provides libnuma
+which provides an interface similar to other NUMA functionality for page
+migration. cat ``/proc/<pid>/numa_maps`` allows an easy review of where the
+pages of a process are located. See also the numa_maps documentation in the
+proc(5) man page.
+
+Manual migration is useful if for example the scheduler has relocated
+a process to a processor on a distant node. A batch scheduler or an
+administrator may detect the situation and move the pages of the process
+nearer to the new processor. The kernel itself only provides
+manual page migration support. Automatic page migration may be implemented
+through user space processes that move pages. A special function call
+"move_pages" allows the moving of individual pages within a process.
+For example, A NUMA profiler may obtain a log showing frequent off-node
+accesses and may use the result to move pages to more advantageous
+locations.
+
+Larger installations usually partition the system using cpusets into
+sections of nodes. Paul Jackson has equipped cpusets with the ability to
+move pages when a task is moved to another cpuset (See
+:ref:`CPUSETS <cpusets>`).
+Cpusets allow the automation of process locality. If a task is moved to
+a new cpuset then also all its pages are moved with it so that the
+performance of the process does not sink dramatically. Also the pages
+of processes in a cpuset are moved if the allowed memory nodes of a
+cpuset are changed.
+
+Page migration allows the preservation of the relative location of pages
+within a group of nodes for all migration techniques which will preserve a
+particular memory allocation pattern generated even after migrating a
+process. This is necessary in order to preserve the memory latencies.
+Processes will run with similar performance after migration.
+
+Page migration occurs in several steps. First a high level
+description for those trying to use migrate_pages() from the kernel
+(for userspace usage see the Andi Kleen's numactl package mentioned above)
+and then a low level description of how the low level details work.
+
+In kernel use of migrate_pages()
+================================
+
+1. Remove pages from the LRU.
+
+ Lists of pages to be migrated are generated by scanning over
+ pages and moving them into lists. This is done by
+ calling isolate_lru_page().
+ Calling isolate_lru_page() increases the references to the page
+ so that it cannot vanish while the page migration occurs.
+ It also prevents the swapper or other scans from encountering
+ the page.
+
+2. We need to have a function of type new_page_t that can be
+ passed to migrate_pages(). This function should figure out
+ how to allocate the correct new page given the old page.
+
+3. The migrate_pages() function is called which attempts
+ to do the migration. It will call the function to allocate
+ the new page for each page that is considered for
+ moving.
+
+How migrate_pages() works
+=========================
+
+migrate_pages() does several passes over its list of pages. A page is moved
+if all references to a page are removable at the time. The page has
+already been removed from the LRU via isolate_lru_page() and the refcount
+is increased so that the page cannot be freed while page migration occurs.
+
+Steps:
+
+1. Lock the page to be migrated.
+
+2. Ensure that writeback is complete.
+
+3. Lock the new page that we want to move to. It is locked so that accesses to
+ this (not yet up-to-date) page immediately block while the move is in progress.
+
+4. All the page table references to the page are converted to migration
+ entries. This decreases the mapcount of a page. If the resulting
+ mapcount is not zero then we do not migrate the page. All user space
+ processes that attempt to access the page will now wait on the page lock
+ or wait for the migration page table entry to be removed.
+
+5. The i_pages lock is taken. This will cause all processes trying
+ to access the page via the mapping to block on the spinlock.
+
+6. The refcount of the page is examined and we back out if references remain.
+ Otherwise, we know that we are the only one referencing this page.
+
+7. The radix tree is checked and if it does not contain the pointer to this
+ page then we back out because someone else modified the radix tree.
+
+8. The new page is prepped with some settings from the old page so that
+ accesses to the new page will discover a page with the correct settings.
+
+9. The radix tree is changed to point to the new page.
+
+10. The reference count of the old page is dropped because the address space
+ reference is gone. A reference to the new page is established because
+ the new page is referenced by the address space.
+
+11. The i_pages lock is dropped. With that lookups in the mapping
+ become possible again. Processes will move from spinning on the lock
+ to sleeping on the locked new page.
+
+12. The page contents are copied to the new page.
+
+13. The remaining page flags are copied to the new page.
+
+14. The old page flags are cleared to indicate that the page does
+ not provide any information anymore.
+
+15. Queued up writeback on the new page is triggered.
+
+16. If migration entries were inserted into the page table, then replace them
+ with real ptes. Doing so will enable access for user space processes not
+ already waiting for the page lock.
+
+17. The page locks are dropped from the old and new page.
+ Processes waiting on the page lock will redo their page faults
+ and will reach the new page.
+
+18. The new page is moved to the LRU and can be scanned by the swapper,
+ etc. again.
+
+Non-LRU page migration
+======================
+
+Although migration originally aimed for reducing the latency of memory accesses
+for NUMA, compaction also uses migration to create high-order pages.
+
+Current problem of the implementation is that it is designed to migrate only
+*LRU* pages. However, there are potential non-LRU pages which can be migrated
+in drivers, for example, zsmalloc, virtio-balloon pages.
+
+For virtio-balloon pages, some parts of migration code path have been hooked
+up and added virtio-balloon specific functions to intercept migration logics.
+It's too specific to a driver so other drivers who want to make their pages
+movable would have to add their own specific hooks in the migration path.
+
+To overcome the problem, VM supports non-LRU page migration which provides
+generic functions for non-LRU movable pages without driver specific hooks
+in the migration path.
+
+If a driver wants to make its pages movable, it should define three functions
+which are function pointers of struct address_space_operations.
+
+1. ``bool (*isolate_page) (struct page *page, isolate_mode_t mode);``
+
+ What VM expects from isolate_page() function of driver is to return *true*
+ if driver isolates the page successfully. On returning true, VM marks the page
+ as PG_isolated so concurrent isolation in several CPUs skip the page
+ for isolation. If a driver cannot isolate the page, it should return *false*.
+
+ Once page is successfully isolated, VM uses page.lru fields so driver
+ shouldn't expect to preserve values in those fields.
+
+2. ``int (*migratepage) (struct address_space *mapping,``
+| ``struct page *newpage, struct page *oldpage, enum migrate_mode);``
+
+ After isolation, VM calls migratepage() of driver with the isolated page.
+ The function of migratepage() is to move the contents of the old page to the
+ new page
+ and set up fields of struct page newpage. Keep in mind that you should
+ indicate to the VM the oldpage is no longer movable via __ClearPageMovable()
+ under page_lock if you migrated the oldpage successfully and returned
+ MIGRATEPAGE_SUCCESS. If driver cannot migrate the page at the moment, driver
+ can return -EAGAIN. On -EAGAIN, VM will retry page migration in a short time
+ because VM interprets -EAGAIN as "temporary migration failure". On returning
+ any error except -EAGAIN, VM will give up the page migration without
+ retrying.
+
+ Driver shouldn't touch the page.lru field while in the migratepage() function.
+
+3. ``void (*putback_page)(struct page *);``
+
+ If migration fails on the isolated page, VM should return the isolated page
+ to the driver so VM calls the driver's putback_page() with the isolated page.
+ In this function, the driver should put the isolated page back into its own data
+ structure.
+
+4. non-LRU movable page flags
+
+ There are two page flags for supporting non-LRU movable page.
+
+ * PG_movable
+
+ Driver should use the function below to make page movable under page_lock::
+
+ void __SetPageMovable(struct page *page, struct address_space *mapping)
+
+ It needs argument of address_space for registering migration
+ family functions which will be called by VM. Exactly speaking,
+ PG_movable is not a real flag of struct page. Rather, VM
+ reuses the page->mapping's lower bits to represent it::
+
+ #define PAGE_MAPPING_MOVABLE 0x2
+ page->mapping = page->mapping | PAGE_MAPPING_MOVABLE;
+
+ so driver shouldn't access page->mapping directly. Instead, driver should
+ use page_mapping() which masks off the low two bits of page->mapping under
+ page lock so it can get the right struct address_space.
+
+ For testing of non-LRU movable pages, VM supports __PageMovable() function.
+ However, it doesn't guarantee to identify non-LRU movable pages because
+ the page->mapping field is unified with other variables in struct page.
+ If the driver releases the page after isolation by VM, page->mapping
+ doesn't have a stable value although it has PAGE_MAPPING_MOVABLE set
+ (look at __ClearPageMovable). But __PageMovable() is cheap to call whether
+ page is LRU or non-LRU movable once the page has been isolated because LRU
+ pages can never have PAGE_MAPPING_MOVABLE set in page->mapping. It is also
+ good for just peeking to test non-LRU movable pages before more expensive
+ checking with lock_page() in pfn scanning to select a victim.
+
+ For guaranteeing non-LRU movable page, VM provides PageMovable() function.
+ Unlike __PageMovable(), PageMovable() validates page->mapping and
+ mapping->a_ops->isolate_page under lock_page(). The lock_page() prevents
+ sudden destroying of page->mapping.
+
+ Drivers using __SetPageMovable() should clear the flag via
+ __ClearMovablePage() under page_lock() before the releasing the page.
+
+ * PG_isolated
+
+ To prevent concurrent isolation among several CPUs, VM marks isolated page
+ as PG_isolated under lock_page(). So if a CPU encounters PG_isolated
+ non-LRU movable page, it can skip it. Driver doesn't need to manipulate the
+ flag because VM will set/clear it automatically. Keep in mind that if the
+ driver sees a PG_isolated page, it means the page has been isolated by the
+ VM so it shouldn't touch the page.lru field.
+ The PG_isolated flag is aliased with the PG_reclaim flag so drivers
+ shouldn't use PG_isolated for its own purposes.
+
+Monitoring Migration
+=====================
+
+The following events (counters) can be used to monitor page migration.
+
+1. PGMIGRATE_SUCCESS: Normal page migration success. Each count means that a
+ page was migrated. If the page was a non-THP page, then this counter is
+ increased by one. If the page was a THP, then this counter is increased by
+ the number of THP subpages. For example, migration of a single 2MB THP that
+ has 4KB-size base pages (subpages) will cause this counter to increase by
+ 512.
+
+2. PGMIGRATE_FAIL: Normal page migration failure. Same counting rules as for
+ PGMIGRATE_SUCCESS, above: this will be increased by the number of subpages,
+ if it was a THP.
+
+3. THP_MIGRATION_SUCCESS: A THP was migrated without being split.
+
+4. THP_MIGRATION_FAIL: A THP could not be migrated nor it could be split.
+
+5. THP_MIGRATION_SPLIT: A THP was migrated, but not as such: first, the THP had
+ to be split. After splitting, a migration retry was used for it's sub-pages.
+
+THP_MIGRATION_* events also update the appropriate PGMIGRATE_SUCCESS or
+PGMIGRATE_FAIL events. For example, a THP migration failure will cause both
+THP_MIGRATION_FAIL and PGMIGRATE_FAIL to increase.
+
+Christoph Lameter, May 8, 2006.
+Minchan Kim, Mar 28, 2016.
diff --git a/Documentation/vm/page_owner.rst b/Documentation/vm/page_owner.rst
new file mode 100644
index 000000000..02deac766
--- /dev/null
+++ b/Documentation/vm/page_owner.rst
@@ -0,0 +1,89 @@
+.. _page_owner:
+
+==================================================
+page owner: Tracking about who allocated each page
+==================================================
+
+Introduction
+============
+
+page owner is for the tracking about who allocated each page.
+It can be used to debug memory leak or to find a memory hogger.
+When allocation happens, information about allocation such as call stack
+and order of pages is stored into certain storage for each page.
+When we need to know about status of all pages, we can get and analyze
+this information.
+
+Although we already have tracepoint for tracing page allocation/free,
+using it for analyzing who allocate each page is rather complex. We need
+to enlarge the trace buffer for preventing overlapping until userspace
+program launched. And, launched program continually dump out the trace
+buffer for later analysis and it would change system behaviour with more
+possibility rather than just keeping it in memory, so bad for debugging.
+
+page owner can also be used for various purposes. For example, accurate
+fragmentation statistics can be obtained through gfp flag information of
+each page. It is already implemented and activated if page owner is
+enabled. Other usages are more than welcome.
+
+page owner is disabled in default. So, if you'd like to use it, you need
+to add "page_owner=on" into your boot cmdline. If the kernel is built
+with page owner and page owner is disabled in runtime due to no enabling
+boot option, runtime overhead is marginal. If disabled in runtime, it
+doesn't require memory to store owner information, so there is no runtime
+memory overhead. And, page owner inserts just two unlikely branches into
+the page allocator hotpath and if not enabled, then allocation is done
+like as the kernel without page owner. These two unlikely branches should
+not affect to allocation performance, especially if the static keys jump
+label patching functionality is available. Following is the kernel's code
+size change due to this facility.
+
+- Without page owner::
+
+ text data bss dec hex filename
+ 40662 1493 644 42799 a72f mm/page_alloc.o
+
+- With page owner::
+
+ text data bss dec hex filename
+ 40892 1493 644 43029 a815 mm/page_alloc.o
+ 1427 24 8 1459 5b3 mm/page_ext.o
+ 2722 50 0 2772 ad4 mm/page_owner.o
+
+Although, roughly, 4 KB code is added in total, page_alloc.o increase by
+230 bytes and only half of it is in hotpath. Building the kernel with
+page owner and turning it on if needed would be great option to debug
+kernel memory problem.
+
+There is one notice that is caused by implementation detail. page owner
+stores information into the memory from struct page extension. This memory
+is initialized some time later than that page allocator starts in sparse
+memory system, so, until initialization, many pages can be allocated and
+they would have no owner information. To fix it up, these early allocated
+pages are investigated and marked as allocated in initialization phase.
+Although it doesn't mean that they have the right owner information,
+at least, we can tell whether the page is allocated or not,
+more accurately. On 2GB memory x86-64 VM box, 13343 early allocated pages
+are catched and marked, although they are mostly allocated from struct
+page extension feature. Anyway, after that, no page is left in
+un-tracking state.
+
+Usage
+=====
+
+1) Build user-space helper::
+
+ cd tools/vm
+ make page_owner_sort
+
+2) Enable page owner: add "page_owner=on" to boot cmdline.
+
+3) Do the job what you want to debug
+
+4) Analyze information from page owner::
+
+ cat /sys/kernel/debug/page_owner > page_owner_full.txt
+ ./page_owner_sort page_owner_full.txt sorted_page_owner.txt
+
+ See the result about who allocated each page
+ in the ``sorted_page_owner.txt``.
diff --git a/Documentation/vm/remap_file_pages.rst b/Documentation/vm/remap_file_pages.rst
new file mode 100644
index 000000000..7bef6718e
--- /dev/null
+++ b/Documentation/vm/remap_file_pages.rst
@@ -0,0 +1,33 @@
+.. _remap_file_pages:
+
+==============================
+remap_file_pages() system call
+==============================
+
+The remap_file_pages() system call is used to create a nonlinear mapping,
+that is, a mapping in which the pages of the file are mapped into a
+nonsequential order in memory. The advantage of using remap_file_pages()
+over using repeated calls to mmap(2) is that the former approach does not
+require the kernel to create additional VMA (Virtual Memory Area) data
+structures.
+
+Supporting of nonlinear mapping requires significant amount of non-trivial
+code in kernel virtual memory subsystem including hot paths. Also to get
+nonlinear mapping work kernel need a way to distinguish normal page table
+entries from entries with file offset (pte_file). Kernel reserves flag in
+PTE for this purpose. PTE flags are scarce resource especially on some CPU
+architectures. It would be nice to free up the flag for other usage.
+
+Fortunately, there are not many users of remap_file_pages() in the wild.
+It's only known that one enterprise RDBMS implementation uses the syscall
+on 32-bit systems to map files bigger than can linearly fit into 32-bit
+virtual address space. This use-case is not critical anymore since 64-bit
+systems are widely available.
+
+The syscall is deprecated and replaced it with an emulation now. The
+emulation creates new VMAs instead of nonlinear mappings. It's going to
+work slower for rare users of remap_file_pages() but ABI is preserved.
+
+One side effect of emulation (apart from performance) is that user can hit
+vm.max_map_count limit more easily due to additional VMAs. See comment for
+DEFAULT_MAX_MAP_COUNT for more details on the limit.
diff --git a/Documentation/vm/slub.rst b/Documentation/vm/slub.rst
new file mode 100644
index 000000000..d3028554b
--- /dev/null
+++ b/Documentation/vm/slub.rst
@@ -0,0 +1,388 @@
+.. _slub:
+
+==========================
+Short users guide for SLUB
+==========================
+
+The basic philosophy of SLUB is very different from SLAB. SLAB
+requires rebuilding the kernel to activate debug options for all
+slab caches. SLUB always includes full debugging but it is off by default.
+SLUB can enable debugging only for selected slabs in order to avoid
+an impact on overall system performance which may make a bug more
+difficult to find.
+
+In order to switch debugging on one can add an option ``slub_debug``
+to the kernel command line. That will enable full debugging for
+all slabs.
+
+Typically one would then use the ``slabinfo`` command to get statistical
+data and perform operation on the slabs. By default ``slabinfo`` only lists
+slabs that have data in them. See "slabinfo -h" for more options when
+running the command. ``slabinfo`` can be compiled with
+::
+
+ gcc -o slabinfo tools/vm/slabinfo.c
+
+Some of the modes of operation of ``slabinfo`` require that slub debugging
+be enabled on the command line. F.e. no tracking information will be
+available without debugging on and validation can only partially
+be performed if debugging was not switched on.
+
+Some more sophisticated uses of slub_debug:
+-------------------------------------------
+
+Parameters may be given to ``slub_debug``. If none is specified then full
+debugging is enabled. Format:
+
+slub_debug=<Debug-Options>
+ Enable options for all slabs
+
+slub_debug=<Debug-Options>,<slab name1>,<slab name2>,...
+ Enable options only for select slabs (no spaces
+ after a comma)
+
+Multiple blocks of options for all slabs or selected slabs can be given, with
+blocks of options delimited by ';'. The last of "all slabs" blocks is applied
+to all slabs except those that match one of the "select slabs" block. Options
+of the first "select slabs" blocks that matches the slab's name are applied.
+
+Possible debug options are::
+
+ F Sanity checks on (enables SLAB_DEBUG_CONSISTENCY_CHECKS
+ Sorry SLAB legacy issues)
+ Z Red zoning
+ P Poisoning (object and padding)
+ U User tracking (free and alloc)
+ T Trace (please only use on single slabs)
+ A Enable failslab filter mark for the cache
+ O Switch debugging off for caches that would have
+ caused higher minimum slab orders
+ - Switch all debugging off (useful if the kernel is
+ configured with CONFIG_SLUB_DEBUG_ON)
+
+F.e. in order to boot just with sanity checks and red zoning one would specify::
+
+ slub_debug=FZ
+
+Trying to find an issue in the dentry cache? Try::
+
+ slub_debug=,dentry
+
+to only enable debugging on the dentry cache. You may use an asterisk at the
+end of the slab name, in order to cover all slabs with the same prefix. For
+example, here's how you can poison the dentry cache as well as all kmalloc
+slabs::
+
+ slub_debug=P,kmalloc-*,dentry
+
+Red zoning and tracking may realign the slab. We can just apply sanity checks
+to the dentry cache with::
+
+ slub_debug=F,dentry
+
+Debugging options may require the minimum possible slab order to increase as
+a result of storing the metadata (for example, caches with PAGE_SIZE object
+sizes). This has a higher liklihood of resulting in slab allocation errors
+in low memory situations or if there's high fragmentation of memory. To
+switch off debugging for such caches by default, use::
+
+ slub_debug=O
+
+You can apply different options to different list of slab names, using blocks
+of options. This will enable red zoning for dentry and user tracking for
+kmalloc. All other slabs will not get any debugging enabled::
+
+ slub_debug=Z,dentry;U,kmalloc-*
+
+You can also enable options (e.g. sanity checks and poisoning) for all caches
+except some that are deemed too performance critical and don't need to be
+debugged by specifying global debug options followed by a list of slab names
+with "-" as options::
+
+ slub_debug=FZ;-,zs_handle,zspage
+
+The state of each debug option for a slab can be found in the respective files
+under::
+
+ /sys/kernel/slab/<slab name>/
+
+If the file contains 1, the option is enabled, 0 means disabled. The debug
+options from the ``slub_debug`` parameter translate to the following files::
+
+ F sanity_checks
+ Z red_zone
+ P poison
+ U store_user
+ T trace
+ A failslab
+
+Careful with tracing: It may spew out lots of information and never stop if
+used on the wrong slab.
+
+Slab merging
+============
+
+If no debug options are specified then SLUB may merge similar slabs together
+in order to reduce overhead and increase cache hotness of objects.
+``slabinfo -a`` displays which slabs were merged together.
+
+Slab validation
+===============
+
+SLUB can validate all object if the kernel was booted with slub_debug. In
+order to do so you must have the ``slabinfo`` tool. Then you can do
+::
+
+ slabinfo -v
+
+which will test all objects. Output will be generated to the syslog.
+
+This also works in a more limited way if boot was without slab debug.
+In that case ``slabinfo -v`` simply tests all reachable objects. Usually
+these are in the cpu slabs and the partial slabs. Full slabs are not
+tracked by SLUB in a non debug situation.
+
+Getting more performance
+========================
+
+To some degree SLUB's performance is limited by the need to take the
+list_lock once in a while to deal with partial slabs. That overhead is
+governed by the order of the allocation for each slab. The allocations
+can be influenced by kernel parameters:
+
+.. slub_min_objects=x (default 4)
+.. slub_min_order=x (default 0)
+.. slub_max_order=x (default 3 (PAGE_ALLOC_COSTLY_ORDER))
+
+``slub_min_objects``
+ allows to specify how many objects must at least fit into one
+ slab in order for the allocation order to be acceptable. In
+ general slub will be able to perform this number of
+ allocations on a slab without consulting centralized resources
+ (list_lock) where contention may occur.
+
+``slub_min_order``
+ specifies a minimum order of slabs. A similar effect like
+ ``slub_min_objects``.
+
+``slub_max_order``
+ specified the order at which ``slub_min_objects`` should no
+ longer be checked. This is useful to avoid SLUB trying to
+ generate super large order pages to fit ``slub_min_objects``
+ of a slab cache with large object sizes into one high order
+ page. Setting command line parameter
+ ``debug_guardpage_minorder=N`` (N > 0), forces setting
+ ``slub_max_order`` to 0, what cause minimum possible order of
+ slabs allocation.
+
+SLUB Debug output
+=================
+
+Here is a sample of slub debug output::
+
+ ====================================================================
+ BUG kmalloc-8: Right Redzone overwritten
+ --------------------------------------------------------------------
+
+ INFO: 0xc90f6d28-0xc90f6d2b. First byte 0x00 instead of 0xcc
+ INFO: Slab 0xc528c530 flags=0x400000c3 inuse=61 fp=0xc90f6d58
+ INFO: Object 0xc90f6d20 @offset=3360 fp=0xc90f6d58
+ INFO: Allocated in get_modalias+0x61/0xf5 age=53 cpu=1 pid=554
+
+ Bytes b4 (0xc90f6d10): 00 00 00 00 00 00 00 00 5a 5a 5a 5a 5a 5a 5a 5a ........ZZZZZZZZ
+ Object (0xc90f6d20): 31 30 31 39 2e 30 30 35 1019.005
+ Redzone (0xc90f6d28): 00 cc cc cc .
+ Padding (0xc90f6d50): 5a 5a 5a 5a 5a 5a 5a 5a ZZZZZZZZ
+
+ [<c010523d>] dump_trace+0x63/0x1eb
+ [<c01053df>] show_trace_log_lvl+0x1a/0x2f
+ [<c010601d>] show_trace+0x12/0x14
+ [<c0106035>] dump_stack+0x16/0x18
+ [<c017e0fa>] object_err+0x143/0x14b
+ [<c017e2cc>] check_object+0x66/0x234
+ [<c017eb43>] __slab_free+0x239/0x384
+ [<c017f446>] kfree+0xa6/0xc6
+ [<c02e2335>] get_modalias+0xb9/0xf5
+ [<c02e23b7>] dmi_dev_uevent+0x27/0x3c
+ [<c027866a>] dev_uevent+0x1ad/0x1da
+ [<c0205024>] kobject_uevent_env+0x20a/0x45b
+ [<c020527f>] kobject_uevent+0xa/0xf
+ [<c02779f1>] store_uevent+0x4f/0x58
+ [<c027758e>] dev_attr_store+0x29/0x2f
+ [<c01bec4f>] sysfs_write_file+0x16e/0x19c
+ [<c0183ba7>] vfs_write+0xd1/0x15a
+ [<c01841d7>] sys_write+0x3d/0x72
+ [<c0104112>] sysenter_past_esp+0x5f/0x99
+ [<b7f7b410>] 0xb7f7b410
+ =======================
+
+ FIX kmalloc-8: Restoring Redzone 0xc90f6d28-0xc90f6d2b=0xcc
+
+If SLUB encounters a corrupted object (full detection requires the kernel
+to be booted with slub_debug) then the following output will be dumped
+into the syslog:
+
+1. Description of the problem encountered
+
+ This will be a message in the system log starting with::
+
+ ===============================================
+ BUG <slab cache affected>: <What went wrong>
+ -----------------------------------------------
+
+ INFO: <corruption start>-<corruption_end> <more info>
+ INFO: Slab <address> <slab information>
+ INFO: Object <address> <object information>
+ INFO: Allocated in <kernel function> age=<jiffies since alloc> cpu=<allocated by
+ cpu> pid=<pid of the process>
+ INFO: Freed in <kernel function> age=<jiffies since free> cpu=<freed by cpu>
+ pid=<pid of the process>
+
+ (Object allocation / free information is only available if SLAB_STORE_USER is
+ set for the slab. slub_debug sets that option)
+
+2. The object contents if an object was involved.
+
+ Various types of lines can follow the BUG SLUB line:
+
+ Bytes b4 <address> : <bytes>
+ Shows a few bytes before the object where the problem was detected.
+ Can be useful if the corruption does not stop with the start of the
+ object.
+
+ Object <address> : <bytes>
+ The bytes of the object. If the object is inactive then the bytes
+ typically contain poison values. Any non-poison value shows a
+ corruption by a write after free.
+
+ Redzone <address> : <bytes>
+ The Redzone following the object. The Redzone is used to detect
+ writes after the object. All bytes should always have the same
+ value. If there is any deviation then it is due to a write after
+ the object boundary.
+
+ (Redzone information is only available if SLAB_RED_ZONE is set.
+ slub_debug sets that option)
+
+ Padding <address> : <bytes>
+ Unused data to fill up the space in order to get the next object
+ properly aligned. In the debug case we make sure that there are
+ at least 4 bytes of padding. This allows the detection of writes
+ before the object.
+
+3. A stackdump
+
+ The stackdump describes the location where the error was detected. The cause
+ of the corruption is may be more likely found by looking at the function that
+ allocated or freed the object.
+
+4. Report on how the problem was dealt with in order to ensure the continued
+ operation of the system.
+
+ These are messages in the system log beginning with::
+
+ FIX <slab cache affected>: <corrective action taken>
+
+ In the above sample SLUB found that the Redzone of an active object has
+ been overwritten. Here a string of 8 characters was written into a slab that
+ has the length of 8 characters. However, a 8 character string needs a
+ terminating 0. That zero has overwritten the first byte of the Redzone field.
+ After reporting the details of the issue encountered the FIX SLUB message
+ tells us that SLUB has restored the Redzone to its proper value and then
+ system operations continue.
+
+Emergency operations
+====================
+
+Minimal debugging (sanity checks alone) can be enabled by booting with::
+
+ slub_debug=F
+
+This will be generally be enough to enable the resiliency features of slub
+which will keep the system running even if a bad kernel component will
+keep corrupting objects. This may be important for production systems.
+Performance will be impacted by the sanity checks and there will be a
+continual stream of error messages to the syslog but no additional memory
+will be used (unlike full debugging).
+
+No guarantees. The kernel component still needs to be fixed. Performance
+may be optimized further by locating the slab that experiences corruption
+and enabling debugging only for that cache
+
+I.e.::
+
+ slub_debug=F,dentry
+
+If the corruption occurs by writing after the end of the object then it
+may be advisable to enable a Redzone to avoid corrupting the beginning
+of other objects::
+
+ slub_debug=FZ,dentry
+
+Extended slabinfo mode and plotting
+===================================
+
+The ``slabinfo`` tool has a special 'extended' ('-X') mode that includes:
+ - Slabcache Totals
+ - Slabs sorted by size (up to -N <num> slabs, default 1)
+ - Slabs sorted by loss (up to -N <num> slabs, default 1)
+
+Additionally, in this mode ``slabinfo`` does not dynamically scale
+sizes (G/M/K) and reports everything in bytes (this functionality is
+also available to other slabinfo modes via '-B' option) which makes
+reporting more precise and accurate. Moreover, in some sense the `-X'
+mode also simplifies the analysis of slabs' behaviour, because its
+output can be plotted using the ``slabinfo-gnuplot.sh`` script. So it
+pushes the analysis from looking through the numbers (tons of numbers)
+to something easier -- visual analysis.
+
+To generate plots:
+
+a) collect slabinfo extended records, for example::
+
+ while [ 1 ]; do slabinfo -X >> FOO_STATS; sleep 1; done
+
+b) pass stats file(-s) to ``slabinfo-gnuplot.sh`` script::
+
+ slabinfo-gnuplot.sh FOO_STATS [FOO_STATS2 .. FOO_STATSN]
+
+ The ``slabinfo-gnuplot.sh`` script will pre-processes the collected records
+ and generates 3 png files (and 3 pre-processing cache files) per STATS
+ file:
+ - Slabcache Totals: FOO_STATS-totals.png
+ - Slabs sorted by size: FOO_STATS-slabs-by-size.png
+ - Slabs sorted by loss: FOO_STATS-slabs-by-loss.png
+
+Another use case, when ``slabinfo-gnuplot.sh`` can be useful, is when you
+need to compare slabs' behaviour "prior to" and "after" some code
+modification. To help you out there, ``slabinfo-gnuplot.sh`` script
+can 'merge' the `Slabcache Totals` sections from different
+measurements. To visually compare N plots:
+
+a) Collect as many STATS1, STATS2, .. STATSN files as you need::
+
+ while [ 1 ]; do slabinfo -X >> STATS<X>; sleep 1; done
+
+b) Pre-process those STATS files::
+
+ slabinfo-gnuplot.sh STATS1 STATS2 .. STATSN
+
+c) Execute ``slabinfo-gnuplot.sh`` in '-t' mode, passing all of the
+ generated pre-processed \*-totals::
+
+ slabinfo-gnuplot.sh -t STATS1-totals STATS2-totals .. STATSN-totals
+
+ This will produce a single plot (png file).
+
+ Plots, expectedly, can be large so some fluctuations or small spikes
+ can go unnoticed. To deal with that, ``slabinfo-gnuplot.sh`` has two
+ options to 'zoom-in'/'zoom-out':
+
+ a) ``-s %d,%d`` -- overwrites the default image width and height
+ b) ``-r %d,%d`` -- specifies a range of samples to use (for example,
+ in ``slabinfo -X >> FOO_STATS; sleep 1;`` case, using a ``-r
+ 40,60`` range will plot only samples collected between 40th and
+ 60th seconds).
+
+Christoph Lameter, May 30, 2007
+Sergey Senozhatsky, October 23, 2015
diff --git a/Documentation/vm/split_page_table_lock.rst b/Documentation/vm/split_page_table_lock.rst
new file mode 100644
index 000000000..ff51f4a54
--- /dev/null
+++ b/Documentation/vm/split_page_table_lock.rst
@@ -0,0 +1,100 @@
+.. _split_page_table_lock:
+
+=====================
+Split page table lock
+=====================
+
+Originally, mm->page_table_lock spinlock protected all page tables of the
+mm_struct. But this approach leads to poor page fault scalability of
+multi-threaded applications due high contention on the lock. To improve
+scalability, split page table lock was introduced.
+
+With split page table lock we have separate per-table lock to serialize
+access to the table. At the moment we use split lock for PTE and PMD
+tables. Access to higher level tables protected by mm->page_table_lock.
+
+There are helpers to lock/unlock a table and other accessor functions:
+
+ - pte_offset_map_lock()
+ maps pte and takes PTE table lock, returns pointer to the taken
+ lock;
+ - pte_unmap_unlock()
+ unlocks and unmaps PTE table;
+ - pte_alloc_map_lock()
+ allocates PTE table if needed and take the lock, returns pointer
+ to taken lock or NULL if allocation failed;
+ - pte_lockptr()
+ returns pointer to PTE table lock;
+ - pmd_lock()
+ takes PMD table lock, returns pointer to taken lock;
+ - pmd_lockptr()
+ returns pointer to PMD table lock;
+
+Split page table lock for PTE tables is enabled compile-time if
+CONFIG_SPLIT_PTLOCK_CPUS (usually 4) is less or equal to NR_CPUS.
+If split lock is disabled, all tables guaded by mm->page_table_lock.
+
+Split page table lock for PMD tables is enabled, if it's enabled for PTE
+tables and the architecture supports it (see below).
+
+Hugetlb and split page table lock
+=================================
+
+Hugetlb can support several page sizes. We use split lock only for PMD
+level, but not for PUD.
+
+Hugetlb-specific helpers:
+
+ - huge_pte_lock()
+ takes pmd split lock for PMD_SIZE page, mm->page_table_lock
+ otherwise;
+ - huge_pte_lockptr()
+ returns pointer to table lock;
+
+Support of split page table lock by an architecture
+===================================================
+
+There's no need in special enabling of PTE split page table lock: everything
+required is done by pgtable_pte_page_ctor() and pgtable_pte_page_dtor(), which
+must be called on PTE table allocation / freeing.
+
+Make sure the architecture doesn't use slab allocator for page table
+allocation: slab uses page->slab_cache for its pages.
+This field shares storage with page->ptl.
+
+PMD split lock only makes sense if you have more than two page table
+levels.
+
+PMD split lock enabling requires pgtable_pmd_page_ctor() call on PMD table
+allocation and pgtable_pmd_page_dtor() on freeing.
+
+Allocation usually happens in pmd_alloc_one(), freeing in pmd_free() and
+pmd_free_tlb(), but make sure you cover all PMD table allocation / freeing
+paths: i.e X86_PAE preallocate few PMDs on pgd_alloc().
+
+With everything in place you can set CONFIG_ARCH_ENABLE_SPLIT_PMD_PTLOCK.
+
+NOTE: pgtable_pte_page_ctor() and pgtable_pmd_page_ctor() can fail -- it must
+be handled properly.
+
+page->ptl
+=========
+
+page->ptl is used to access split page table lock, where 'page' is struct
+page of page containing the table. It shares storage with page->private
+(and few other fields in union).
+
+To avoid increasing size of struct page and have best performance, we use a
+trick:
+
+ - if spinlock_t fits into long, we use page->ptr as spinlock, so we
+ can avoid indirect access and save a cache line.
+ - if size of spinlock_t is bigger then size of long, we use page->ptl as
+ pointer to spinlock_t and allocate it dynamically. This allows to use
+ split lock with enabled DEBUG_SPINLOCK or DEBUG_LOCK_ALLOC, but costs
+ one more cache line for indirect access;
+
+The spinlock_t allocated in pgtable_pte_page_ctor() for PTE table and in
+pgtable_pmd_page_ctor() for PMD table.
+
+Please, never access page->ptl directly -- use appropriate helper.
diff --git a/Documentation/vm/swap_numa.rst b/Documentation/vm/swap_numa.rst
new file mode 100644
index 000000000..e0466f2db
--- /dev/null
+++ b/Documentation/vm/swap_numa.rst
@@ -0,0 +1,80 @@
+.. _swap_numa:
+
+===========================================
+Automatically bind swap device to numa node
+===========================================
+
+If the system has more than one swap device and swap device has the node
+information, we can make use of this information to decide which swap
+device to use in get_swap_pages() to get better performance.
+
+
+How to use this feature
+=======================
+
+Swap device has priority and that decides the order of it to be used. To make
+use of automatically binding, there is no need to manipulate priority settings
+for swap devices. e.g. on a 2 node machine, assume 2 swap devices swapA and
+swapB, with swapA attached to node 0 and swapB attached to node 1, are going
+to be swapped on. Simply swapping them on by doing::
+
+ # swapon /dev/swapA
+ # swapon /dev/swapB
+
+Then node 0 will use the two swap devices in the order of swapA then swapB and
+node 1 will use the two swap devices in the order of swapB then swapA. Note
+that the order of them being swapped on doesn't matter.
+
+A more complex example on a 4 node machine. Assume 6 swap devices are going to
+be swapped on: swapA and swapB are attached to node 0, swapC is attached to
+node 1, swapD and swapE are attached to node 2 and swapF is attached to node3.
+The way to swap them on is the same as above::
+
+ # swapon /dev/swapA
+ # swapon /dev/swapB
+ # swapon /dev/swapC
+ # swapon /dev/swapD
+ # swapon /dev/swapE
+ # swapon /dev/swapF
+
+Then node 0 will use them in the order of::
+
+ swapA/swapB -> swapC -> swapD -> swapE -> swapF
+
+swapA and swapB will be used in a round robin mode before any other swap device.
+
+node 1 will use them in the order of::
+
+ swapC -> swapA -> swapB -> swapD -> swapE -> swapF
+
+node 2 will use them in the order of::
+
+ swapD/swapE -> swapA -> swapB -> swapC -> swapF
+
+Similaly, swapD and swapE will be used in a round robin mode before any
+other swap devices.
+
+node 3 will use them in the order of::
+
+ swapF -> swapA -> swapB -> swapC -> swapD -> swapE
+
+
+Implementation details
+======================
+
+The current code uses a priority based list, swap_avail_list, to decide
+which swap device to use and if multiple swap devices share the same
+priority, they are used round robin. This change here replaces the single
+global swap_avail_list with a per-numa-node list, i.e. for each numa node,
+it sees its own priority based list of available swap devices. Swap
+device's priority can be promoted on its matching node's swap_avail_list.
+
+The current swap device's priority is set as: user can set a >=0 value,
+or the system will pick one starting from -1 then downwards. The priority
+value in the swap_avail_list is the negated value of the swap device's
+due to plist being sorted from low to high. The new policy doesn't change
+the semantics for priority >=0 cases, the previous starting from -1 then
+downwards now becomes starting from -2 then downwards and -1 is reserved
+as the promoted value. So if multiple swap devices are attached to the same
+node, they will all be promoted to priority -1 on that node's plist and will
+be used round robin before any other swap devices.
diff --git a/Documentation/vm/transhuge.rst b/Documentation/vm/transhuge.rst
new file mode 100644
index 000000000..0ed23e59a
--- /dev/null
+++ b/Documentation/vm/transhuge.rst
@@ -0,0 +1,192 @@
+.. _transhuge:
+
+============================
+Transparent Hugepage Support
+============================
+
+This document describes design principles for Transparent Hugepage (THP)
+support and its interaction with other parts of the memory management
+system.
+
+Design principles
+=================
+
+- "graceful fallback": mm components which don't have transparent hugepage
+ knowledge fall back to breaking huge pmd mapping into table of ptes and,
+ if necessary, split a transparent hugepage. Therefore these components
+ can continue working on the regular pages or regular pte mappings.
+
+- if a hugepage allocation fails because of memory fragmentation,
+ regular pages should be gracefully allocated instead and mixed in
+ the same vma without any failure or significant delay and without
+ userland noticing
+
+- if some task quits and more hugepages become available (either
+ immediately in the buddy or through the VM), guest physical memory
+ backed by regular pages should be relocated on hugepages
+ automatically (with khugepaged)
+
+- it doesn't require memory reservation and in turn it uses hugepages
+ whenever possible (the only possible reservation here is kernelcore=
+ to avoid unmovable pages to fragment all the memory but such a tweak
+ is not specific to transparent hugepage support and it's a generic
+ feature that applies to all dynamic high order allocations in the
+ kernel)
+
+get_user_pages and follow_page
+==============================
+
+get_user_pages and follow_page if run on a hugepage, will return the
+head or tail pages as usual (exactly as they would do on
+hugetlbfs). Most GUP users will only care about the actual physical
+address of the page and its temporary pinning to release after the I/O
+is complete, so they won't ever notice the fact the page is huge. But
+if any driver is going to mangle over the page structure of the tail
+page (like for checking page->mapping or other bits that are relevant
+for the head page and not the tail page), it should be updated to jump
+to check head page instead. Taking a reference on any head/tail page would
+prevent the page from being split by anyone.
+
+.. note::
+ these aren't new constraints to the GUP API, and they match the
+ same constraints that apply to hugetlbfs too, so any driver capable
+ of handling GUP on hugetlbfs will also work fine on transparent
+ hugepage backed mappings.
+
+In case you can't handle compound pages if they're returned by
+follow_page, the FOLL_SPLIT bit can be specified as a parameter to
+follow_page, so that it will split the hugepages before returning
+them.
+
+Graceful fallback
+=================
+
+Code walking pagetables but unaware about huge pmds can simply call
+split_huge_pmd(vma, pmd, addr) where the pmd is the one returned by
+pmd_offset. It's trivial to make the code transparent hugepage aware
+by just grepping for "pmd_offset" and adding split_huge_pmd where
+missing after pmd_offset returns the pmd. Thanks to the graceful
+fallback design, with a one liner change, you can avoid to write
+hundreds if not thousands of lines of complex code to make your code
+hugepage aware.
+
+If you're not walking pagetables but you run into a physical hugepage
+that you can't handle natively in your code, you can split it by
+calling split_huge_page(page). This is what the Linux VM does before
+it tries to swapout the hugepage for example. split_huge_page() can fail
+if the page is pinned and you must handle this correctly.
+
+Example to make mremap.c transparent hugepage aware with a one liner
+change::
+
+ diff --git a/mm/mremap.c b/mm/mremap.c
+ --- a/mm/mremap.c
+ +++ b/mm/mremap.c
+ @@ -41,6 +41,7 @@ static pmd_t *get_old_pmd(struct mm_stru
+ return NULL;
+
+ pmd = pmd_offset(pud, addr);
+ + split_huge_pmd(vma, pmd, addr);
+ if (pmd_none_or_clear_bad(pmd))
+ return NULL;
+
+Locking in hugepage aware code
+==============================
+
+We want as much code as possible hugepage aware, as calling
+split_huge_page() or split_huge_pmd() has a cost.
+
+To make pagetable walks huge pmd aware, all you need to do is to call
+pmd_trans_huge() on the pmd returned by pmd_offset. You must hold the
+mmap_lock in read (or write) mode to be sure a huge pmd cannot be
+created from under you by khugepaged (khugepaged collapse_huge_page
+takes the mmap_lock in write mode in addition to the anon_vma lock). If
+pmd_trans_huge returns false, you just fallback in the old code
+paths. If instead pmd_trans_huge returns true, you have to take the
+page table lock (pmd_lock()) and re-run pmd_trans_huge. Taking the
+page table lock will prevent the huge pmd being converted into a
+regular pmd from under you (split_huge_pmd can run in parallel to the
+pagetable walk). If the second pmd_trans_huge returns false, you
+should just drop the page table lock and fallback to the old code as
+before. Otherwise, you can proceed to process the huge pmd and the
+hugepage natively. Once finished, you can drop the page table lock.
+
+Refcounts and transparent huge pages
+====================================
+
+Refcounting on THP is mostly consistent with refcounting on other compound
+pages:
+
+ - get_page()/put_page() and GUP operate on head page's ->_refcount.
+
+ - ->_refcount in tail pages is always zero: get_page_unless_zero() never
+ succeeds on tail pages.
+
+ - map/unmap of the pages with PTE entry increment/decrement ->_mapcount
+ on relevant sub-page of the compound page.
+
+ - map/unmap of the whole compound page is accounted for in compound_mapcount
+ (stored in first tail page). For file huge pages, we also increment
+ ->_mapcount of all sub-pages in order to have race-free detection of
+ last unmap of subpages.
+
+PageDoubleMap() indicates that the page is *possibly* mapped with PTEs.
+
+For anonymous pages, PageDoubleMap() also indicates ->_mapcount in all
+subpages is offset up by one. This additional reference is required to
+get race-free detection of unmap of subpages when we have them mapped with
+both PMDs and PTEs.
+
+This optimization is required to lower the overhead of per-subpage mapcount
+tracking. The alternative is to alter ->_mapcount in all subpages on each
+map/unmap of the whole compound page.
+
+For anonymous pages, we set PG_double_map when a PMD of the page is split
+for the first time, but still have a PMD mapping. The additional references
+go away with the last compound_mapcount.
+
+File pages get PG_double_map set on the first map of the page with PTE and
+goes away when the page gets evicted from the page cache.
+
+split_huge_page internally has to distribute the refcounts in the head
+page to the tail pages before clearing all PG_head/tail bits from the page
+structures. It can be done easily for refcounts taken by page table
+entries, but we don't have enough information on how to distribute any
+additional pins (i.e. from get_user_pages). split_huge_page() fails any
+requests to split pinned huge pages: it expects page count to be equal to
+the sum of mapcount of all sub-pages plus one (split_huge_page caller must
+have a reference to the head page).
+
+split_huge_page uses migration entries to stabilize page->_refcount and
+page->_mapcount of anonymous pages. File pages just get unmapped.
+
+We are safe against physical memory scanners too: the only legitimate way
+a scanner can get a reference to a page is get_page_unless_zero().
+
+All tail pages have zero ->_refcount until atomic_add(). This prevents the
+scanner from getting a reference to the tail page up to that point. After the
+atomic_add() we don't care about the ->_refcount value. We already know how
+many references should be uncharged from the head page.
+
+For head page get_page_unless_zero() will succeed and we don't mind. It's
+clear where references should go after split: it will stay on the head page.
+
+Note that split_huge_pmd() doesn't have any limitations on refcounting:
+pmd can be split at any point and never fails.
+
+Partial unmap and deferred_split_huge_page()
+============================================
+
+Unmapping part of THP (with munmap() or other way) is not going to free
+memory immediately. Instead, we detect that a subpage of THP is not in use
+in page_remove_rmap() and queue the THP for splitting if memory pressure
+comes. Splitting will free up unused subpages.
+
+Splitting the page right away is not an option due to locking context in
+the place where we can detect partial unmap. It also might be
+counterproductive since in many cases partial unmap happens during exit(2) if
+a THP crosses a VMA boundary.
+
+The function deferred_split_huge_page() is used to queue a page for splitting.
+The splitting itself will happen when we get memory pressure via shrinker
+interface.
diff --git a/Documentation/vm/unevictable-lru.rst b/Documentation/vm/unevictable-lru.rst
new file mode 100644
index 000000000..17d0861b0
--- /dev/null
+++ b/Documentation/vm/unevictable-lru.rst
@@ -0,0 +1,618 @@
+.. _unevictable_lru:
+
+==============================
+Unevictable LRU Infrastructure
+==============================
+
+.. contents:: :local:
+
+
+Introduction
+============
+
+This document describes the Linux memory manager's "Unevictable LRU"
+infrastructure and the use of this to manage several types of "unevictable"
+pages.
+
+The document attempts to provide the overall rationale behind this mechanism
+and the rationale for some of the design decisions that drove the
+implementation. The latter design rationale is discussed in the context of an
+implementation description. Admittedly, one can obtain the implementation
+details - the "what does it do?" - by reading the code. One hopes that the
+descriptions below add value by provide the answer to "why does it do that?".
+
+
+
+The Unevictable LRU
+===================
+
+The Unevictable LRU facility adds an additional LRU list to track unevictable
+pages and to hide these pages from vmscan. This mechanism is based on a patch
+by Larry Woodman of Red Hat to address several scalability problems with page
+reclaim in Linux. The problems have been observed at customer sites on large
+memory x86_64 systems.
+
+To illustrate this with an example, a non-NUMA x86_64 platform with 128GB of
+main memory will have over 32 million 4k pages in a single zone. When a large
+fraction of these pages are not evictable for any reason [see below], vmscan
+will spend a lot of time scanning the LRU lists looking for the small fraction
+of pages that are evictable. This can result in a situation where all CPUs are
+spending 100% of their time in vmscan for hours or days on end, with the system
+completely unresponsive.
+
+The unevictable list addresses the following classes of unevictable pages:
+
+ * Those owned by ramfs.
+
+ * Those mapped into SHM_LOCK'd shared memory regions.
+
+ * Those mapped into VM_LOCKED [mlock()ed] VMAs.
+
+The infrastructure may also be able to handle other conditions that make pages
+unevictable, either by definition or by circumstance, in the future.
+
+
+The Unevictable Page List
+-------------------------
+
+The Unevictable LRU infrastructure consists of an additional, per-zone, LRU list
+called the "unevictable" list and an associated page flag, PG_unevictable, to
+indicate that the page is being managed on the unevictable list.
+
+The PG_unevictable flag is analogous to, and mutually exclusive with, the
+PG_active flag in that it indicates on which LRU list a page resides when
+PG_lru is set.
+
+The Unevictable LRU infrastructure maintains unevictable pages on an additional
+LRU list for a few reasons:
+
+ (1) We get to "treat unevictable pages just like we treat other pages in the
+ system - which means we get to use the same code to manipulate them, the
+ same code to isolate them (for migrate, etc.), the same code to keep track
+ of the statistics, etc..." [Rik van Riel]
+
+ (2) We want to be able to migrate unevictable pages between nodes for memory
+ defragmentation, workload management and memory hotplug. The linux kernel
+ can only migrate pages that it can successfully isolate from the LRU
+ lists. If we were to maintain pages elsewhere than on an LRU-like list,
+ where they can be found by isolate_lru_page(), we would prevent their
+ migration, unless we reworked migration code to find the unevictable pages
+ itself.
+
+
+The unevictable list does not differentiate between file-backed and anonymous,
+swap-backed pages. This differentiation is only important while the pages are,
+in fact, evictable.
+
+The unevictable list benefits from the "arrayification" of the per-zone LRU
+lists and statistics originally proposed and posted by Christoph Lameter.
+
+The unevictable list does not use the LRU pagevec mechanism. Rather,
+unevictable pages are placed directly on the page's zone's unevictable list
+under the zone lru_lock. This allows us to prevent the stranding of pages on
+the unevictable list when one task has the page isolated from the LRU and other
+tasks are changing the "evictability" state of the page.
+
+
+Memory Control Group Interaction
+--------------------------------
+
+The unevictable LRU facility interacts with the memory control group [aka
+memory controller; see Documentation/admin-guide/cgroup-v1/memory.rst] by extending the
+lru_list enum.
+
+The memory controller data structure automatically gets a per-zone unevictable
+list as a result of the "arrayification" of the per-zone LRU lists (one per
+lru_list enum element). The memory controller tracks the movement of pages to
+and from the unevictable list.
+
+When a memory control group comes under memory pressure, the controller will
+not attempt to reclaim pages on the unevictable list. This has a couple of
+effects:
+
+ (1) Because the pages are "hidden" from reclaim on the unevictable list, the
+ reclaim process can be more efficient, dealing only with pages that have a
+ chance of being reclaimed.
+
+ (2) On the other hand, if too many of the pages charged to the control group
+ are unevictable, the evictable portion of the working set of the tasks in
+ the control group may not fit into the available memory. This can cause
+ the control group to thrash or to OOM-kill tasks.
+
+
+.. _mark_addr_space_unevict:
+
+Marking Address Spaces Unevictable
+----------------------------------
+
+For facilities such as ramfs none of the pages attached to the address space
+may be evicted. To prevent eviction of any such pages, the AS_UNEVICTABLE
+address space flag is provided, and this can be manipulated by a filesystem
+using a number of wrapper functions:
+
+ * ``void mapping_set_unevictable(struct address_space *mapping);``
+
+ Mark the address space as being completely unevictable.
+
+ * ``void mapping_clear_unevictable(struct address_space *mapping);``
+
+ Mark the address space as being evictable.
+
+ * ``int mapping_unevictable(struct address_space *mapping);``
+
+ Query the address space, and return true if it is completely
+ unevictable.
+
+These are currently used in three places in the kernel:
+
+ (1) By ramfs to mark the address spaces of its inodes when they are created,
+ and this mark remains for the life of the inode.
+
+ (2) By SYSV SHM to mark SHM_LOCK'd address spaces until SHM_UNLOCK is called.
+
+ Note that SHM_LOCK is not required to page in the locked pages if they're
+ swapped out; the application must touch the pages manually if it wants to
+ ensure they're in memory.
+
+ (3) By the i915 driver to mark pinned address space until it's unpinned. The
+ amount of unevictable memory marked by i915 driver is roughly the bounded
+ object size in debugfs/dri/0/i915_gem_objects.
+
+
+Detecting Unevictable Pages
+---------------------------
+
+The function page_evictable() in vmscan.c determines whether a page is
+evictable or not using the query function outlined above [see section
+:ref:`Marking address spaces unevictable <mark_addr_space_unevict>`]
+to check the AS_UNEVICTABLE flag.
+
+For address spaces that are so marked after being populated (as SHM regions
+might be), the lock action (eg: SHM_LOCK) can be lazy, and need not populate
+the page tables for the region as does, for example, mlock(), nor need it make
+any special effort to push any pages in the SHM_LOCK'd area to the unevictable
+list. Instead, vmscan will do this if and when it encounters the pages during
+a reclamation scan.
+
+On an unlock action (such as SHM_UNLOCK), the unlocker (eg: shmctl()) must scan
+the pages in the region and "rescue" them from the unevictable list if no other
+condition is keeping them unevictable. If an unevictable region is destroyed,
+the pages are also "rescued" from the unevictable list in the process of
+freeing them.
+
+page_evictable() also checks for mlocked pages by testing an additional page
+flag, PG_mlocked (as wrapped by PageMlocked()), which is set when a page is
+faulted into a VM_LOCKED vma, or found in a vma being VM_LOCKED.
+
+
+Vmscan's Handling of Unevictable Pages
+--------------------------------------
+
+If unevictable pages are culled in the fault path, or moved to the unevictable
+list at mlock() or mmap() time, vmscan will not encounter the pages until they
+have become evictable again (via munlock() for example) and have been "rescued"
+from the unevictable list. However, there may be situations where we decide,
+for the sake of expediency, to leave a unevictable page on one of the regular
+active/inactive LRU lists for vmscan to deal with. vmscan checks for such
+pages in all of the shrink_{active|inactive|page}_list() functions and will
+"cull" such pages that it encounters: that is, it diverts those pages to the
+unevictable list for the zone being scanned.
+
+There may be situations where a page is mapped into a VM_LOCKED VMA, but the
+page is not marked as PG_mlocked. Such pages will make it all the way to
+shrink_page_list() where they will be detected when vmscan walks the reverse
+map in try_to_unmap(). If try_to_unmap() returns SWAP_MLOCK,
+shrink_page_list() will cull the page at that point.
+
+To "cull" an unevictable page, vmscan simply puts the page back on the LRU list
+using putback_lru_page() - the inverse operation to isolate_lru_page() - after
+dropping the page lock. Because the condition which makes the page unevictable
+may change once the page is unlocked, putback_lru_page() will recheck the
+unevictable state of a page that it places on the unevictable list. If the
+page has become unevictable, putback_lru_page() removes it from the list and
+retries, including the page_unevictable() test. Because such a race is a rare
+event and movement of pages onto the unevictable list should be rare, these
+extra evictabilty checks should not occur in the majority of calls to
+putback_lru_page().
+
+
+MLOCKED Pages
+=============
+
+The unevictable page list is also useful for mlock(), in addition to ramfs and
+SYSV SHM. Note that mlock() is only available in CONFIG_MMU=y situations; in
+NOMMU situations, all mappings are effectively mlocked.
+
+
+History
+-------
+
+The "Unevictable mlocked Pages" infrastructure is based on work originally
+posted by Nick Piggin in an RFC patch entitled "mm: mlocked pages off LRU".
+Nick posted his patch as an alternative to a patch posted by Christoph Lameter
+to achieve the same objective: hiding mlocked pages from vmscan.
+
+In Nick's patch, he used one of the struct page LRU list link fields as a count
+of VM_LOCKED VMAs that map the page. This use of the link field for a count
+prevented the management of the pages on an LRU list, and thus mlocked pages
+were not migratable as isolate_lru_page() could not find them, and the LRU list
+link field was not available to the migration subsystem.
+
+Nick resolved this by putting mlocked pages back on the lru list before
+attempting to isolate them, thus abandoning the count of VM_LOCKED VMAs. When
+Nick's patch was integrated with the Unevictable LRU work, the count was
+replaced by walking the reverse map to determine whether any VM_LOCKED VMAs
+mapped the page. More on this below.
+
+
+Basic Management
+----------------
+
+mlocked pages - pages mapped into a VM_LOCKED VMA - are a class of unevictable
+pages. When such a page has been "noticed" by the memory management subsystem,
+the page is marked with the PG_mlocked flag. This can be manipulated using the
+PageMlocked() functions.
+
+A PG_mlocked page will be placed on the unevictable list when it is added to
+the LRU. Such pages can be "noticed" by memory management in several places:
+
+ (1) in the mlock()/mlockall() system call handlers;
+
+ (2) in the mmap() system call handler when mmapping a region with the
+ MAP_LOCKED flag;
+
+ (3) mmapping a region in a task that has called mlockall() with the MCL_FUTURE
+ flag
+
+ (4) in the fault path, if mlocked pages are "culled" in the fault path,
+ and when a VM_LOCKED stack segment is expanded; or
+
+ (5) as mentioned above, in vmscan:shrink_page_list() when attempting to
+ reclaim a page in a VM_LOCKED VMA via try_to_unmap()
+
+all of which result in the VM_LOCKED flag being set for the VMA if it doesn't
+already have it set.
+
+mlocked pages become unlocked and rescued from the unevictable list when:
+
+ (1) mapped in a range unlocked via the munlock()/munlockall() system calls;
+
+ (2) munmap()'d out of the last VM_LOCKED VMA that maps the page, including
+ unmapping at task exit;
+
+ (3) when the page is truncated from the last VM_LOCKED VMA of an mmapped file;
+ or
+
+ (4) before a page is COW'd in a VM_LOCKED VMA.
+
+
+mlock()/mlockall() System Call Handling
+---------------------------------------
+
+Both [do\_]mlock() and [do\_]mlockall() system call handlers call mlock_fixup()
+for each VMA in the range specified by the call. In the case of mlockall(),
+this is the entire active address space of the task. Note that mlock_fixup()
+is used for both mlocking and munlocking a range of memory. A call to mlock()
+an already VM_LOCKED VMA, or to munlock() a VMA that is not VM_LOCKED is
+treated as a no-op, and mlock_fixup() simply returns.
+
+If the VMA passes some filtering as described in "Filtering Special Vmas"
+below, mlock_fixup() will attempt to merge the VMA with its neighbors or split
+off a subset of the VMA if the range does not cover the entire VMA. Once the
+VMA has been merged or split or neither, mlock_fixup() will call
+populate_vma_page_range() to fault in the pages via get_user_pages() and to
+mark the pages as mlocked via mlock_vma_page().
+
+Note that the VMA being mlocked might be mapped with PROT_NONE. In this case,
+get_user_pages() will be unable to fault in the pages. That's okay. If pages
+do end up getting faulted into this VM_LOCKED VMA, we'll handle them in the
+fault path or in vmscan.
+
+Also note that a page returned by get_user_pages() could be truncated or
+migrated out from under us, while we're trying to mlock it. To detect this,
+populate_vma_page_range() checks page_mapping() after acquiring the page lock.
+If the page is still associated with its mapping, we'll go ahead and call
+mlock_vma_page(). If the mapping is gone, we just unlock the page and move on.
+In the worst case, this will result in a page mapped in a VM_LOCKED VMA
+remaining on a normal LRU list without being PageMlocked(). Again, vmscan will
+detect and cull such pages.
+
+mlock_vma_page() will call TestSetPageMlocked() for each page returned by
+get_user_pages(). We use TestSetPageMlocked() because the page might already
+be mlocked by another task/VMA and we don't want to do extra work. We
+especially do not want to count an mlocked page more than once in the
+statistics. If the page was already mlocked, mlock_vma_page() need do nothing
+more.
+
+If the page was NOT already mlocked, mlock_vma_page() attempts to isolate the
+page from the LRU, as it is likely on the appropriate active or inactive list
+at that time. If the isolate_lru_page() succeeds, mlock_vma_page() will put
+back the page - by calling putback_lru_page() - which will notice that the page
+is now mlocked and divert the page to the zone's unevictable list. If
+mlock_vma_page() is unable to isolate the page from the LRU, vmscan will handle
+it later if and when it attempts to reclaim the page.
+
+
+Filtering Special VMAs
+----------------------
+
+mlock_fixup() filters several classes of "special" VMAs:
+
+1) VMAs with VM_IO or VM_PFNMAP set are skipped entirely. The pages behind
+ these mappings are inherently pinned, so we don't need to mark them as
+ mlocked. In any case, most of the pages have no struct page in which to so
+ mark the page. Because of this, get_user_pages() will fail for these VMAs,
+ so there is no sense in attempting to visit them.
+
+2) VMAs mapping hugetlbfs page are already effectively pinned into memory. We
+ neither need nor want to mlock() these pages. However, to preserve the
+ prior behavior of mlock() - before the unevictable/mlock changes -
+ mlock_fixup() will call make_pages_present() in the hugetlbfs VMA range to
+ allocate the huge pages and populate the ptes.
+
+3) VMAs with VM_DONTEXPAND are generally userspace mappings of kernel pages,
+ such as the VDSO page, relay channel pages, etc. These pages
+ are inherently unevictable and are not managed on the LRU lists.
+ mlock_fixup() treats these VMAs the same as hugetlbfs VMAs. It calls
+ make_pages_present() to populate the ptes.
+
+Note that for all of these special VMAs, mlock_fixup() does not set the
+VM_LOCKED flag. Therefore, we won't have to deal with them later during
+munlock(), munmap() or task exit. Neither does mlock_fixup() account these
+VMAs against the task's "locked_vm".
+
+.. _munlock_munlockall_handling:
+
+munlock()/munlockall() System Call Handling
+-------------------------------------------
+
+The munlock() and munlockall() system calls are handled by the same functions -
+do_mlock[all]() - as the mlock() and mlockall() system calls with the unlock vs
+lock operation indicated by an argument. So, these system calls are also
+handled by mlock_fixup(). Again, if called for an already munlocked VMA,
+mlock_fixup() simply returns. Because of the VMA filtering discussed above,
+VM_LOCKED will not be set in any "special" VMAs. So, these VMAs will be
+ignored for munlock.
+
+If the VMA is VM_LOCKED, mlock_fixup() again attempts to merge or split off the
+specified range. The range is then munlocked via the function
+populate_vma_page_range() - the same function used to mlock a VMA range -
+passing a flag to indicate that munlock() is being performed.
+
+Because the VMA access protections could have been changed to PROT_NONE after
+faulting in and mlocking pages, get_user_pages() was unreliable for visiting
+these pages for munlocking. Because we don't want to leave pages mlocked,
+get_user_pages() was enhanced to accept a flag to ignore the permissions when
+fetching the pages - all of which should be resident as a result of previous
+mlocking.
+
+For munlock(), populate_vma_page_range() unlocks individual pages by calling
+munlock_vma_page(). munlock_vma_page() unconditionally clears the PG_mlocked
+flag using TestClearPageMlocked(). As with mlock_vma_page(),
+munlock_vma_page() use the Test*PageMlocked() function to handle the case where
+the page might have already been unlocked by another task. If the page was
+mlocked, munlock_vma_page() updates that zone statistics for the number of
+mlocked pages. Note, however, that at this point we haven't checked whether
+the page is mapped by other VM_LOCKED VMAs.
+
+We can't call try_to_munlock(), the function that walks the reverse map to
+check for other VM_LOCKED VMAs, without first isolating the page from the LRU.
+try_to_munlock() is a variant of try_to_unmap() and thus requires that the page
+not be on an LRU list [more on these below]. However, the call to
+isolate_lru_page() could fail, in which case we couldn't try_to_munlock(). So,
+we go ahead and clear PG_mlocked up front, as this might be the only chance we
+have. If we can successfully isolate the page, we go ahead and
+try_to_munlock(), which will restore the PG_mlocked flag and update the zone
+page statistics if it finds another VMA holding the page mlocked. If we fail
+to isolate the page, we'll have left a potentially mlocked page on the LRU.
+This is fine, because we'll catch it later if and if vmscan tries to reclaim
+the page. This should be relatively rare.
+
+
+Migrating MLOCKED Pages
+-----------------------
+
+A page that is being migrated has been isolated from the LRU lists and is held
+locked across unmapping of the page, updating the page's address space entry
+and copying the contents and state, until the page table entry has been
+replaced with an entry that refers to the new page. Linux supports migration
+of mlocked pages and other unevictable pages. This involves simply moving the
+PG_mlocked and PG_unevictable states from the old page to the new page.
+
+Note that page migration can race with mlocking or munlocking of the same page.
+This has been discussed from the mlock/munlock perspective in the respective
+sections above. Both processes (migration and m[un]locking) hold the page
+locked. This provides the first level of synchronization. Page migration
+zeros out the page_mapping of the old page before unlocking it, so m[un]lock
+can skip these pages by testing the page mapping under page lock.
+
+To complete page migration, we place the new and old pages back onto the LRU
+after dropping the page lock. The "unneeded" page - old page on success, new
+page on failure - will be freed when the reference count held by the migration
+process is released. To ensure that we don't strand pages on the unevictable
+list because of a race between munlock and migration, page migration uses the
+putback_lru_page() function to add migrated pages back to the LRU.
+
+
+Compacting MLOCKED Pages
+------------------------
+
+The unevictable LRU can be scanned for compactable regions and the default
+behavior is to do so. /proc/sys/vm/compact_unevictable_allowed controls
+this behavior (see Documentation/admin-guide/sysctl/vm.rst). Once scanning of the
+unevictable LRU is enabled, the work of compaction is mostly handled by
+the page migration code and the same work flow as described in MIGRATING
+MLOCKED PAGES will apply.
+
+MLOCKING Transparent Huge Pages
+-------------------------------
+
+A transparent huge page is represented by a single entry on an LRU list.
+Therefore, we can only make unevictable an entire compound page, not
+individual subpages.
+
+If a user tries to mlock() part of a huge page, we want the rest of the
+page to be reclaimable.
+
+We cannot just split the page on partial mlock() as split_huge_page() can
+fail and new intermittent failure mode for the syscall is undesirable.
+
+We handle this by keeping PTE-mapped huge pages on normal LRU lists: the
+PMD on border of VM_LOCKED VMA will be split into PTE table.
+
+This way the huge page is accessible for vmscan. Under memory pressure the
+page will be split, subpages which belong to VM_LOCKED VMAs will be moved
+to unevictable LRU and the rest can be reclaimed.
+
+See also comment in follow_trans_huge_pmd().
+
+mmap(MAP_LOCKED) System Call Handling
+-------------------------------------
+
+In addition the mlock()/mlockall() system calls, an application can request
+that a region of memory be mlocked supplying the MAP_LOCKED flag to the mmap()
+call. There is one important and subtle difference here, though. mmap() + mlock()
+will fail if the range cannot be faulted in (e.g. because mm_populate fails)
+and returns with ENOMEM while mmap(MAP_LOCKED) will not fail. The mmaped
+area will still have properties of the locked area - aka. pages will not get
+swapped out - but major page faults to fault memory in might still happen.
+
+Furthermore, any mmap() call or brk() call that expands the heap by a
+task that has previously called mlockall() with the MCL_FUTURE flag will result
+in the newly mapped memory being mlocked. Before the unevictable/mlock
+changes, the kernel simply called make_pages_present() to allocate pages and
+populate the page table.
+
+To mlock a range of memory under the unevictable/mlock infrastructure, the
+mmap() handler and task address space expansion functions call
+populate_vma_page_range() specifying the vma and the address range to mlock.
+
+The callers of populate_vma_page_range() will have already added the memory range
+to be mlocked to the task's "locked_vm". To account for filtered VMAs,
+populate_vma_page_range() returns the number of pages NOT mlocked. All of the
+callers then subtract a non-negative return value from the task's locked_vm. A
+negative return value represent an error - for example, from get_user_pages()
+attempting to fault in a VMA with PROT_NONE access. In this case, we leave the
+memory range accounted as locked_vm, as the protections could be changed later
+and pages allocated into that region.
+
+
+munmap()/exit()/exec() System Call Handling
+-------------------------------------------
+
+When unmapping an mlocked region of memory, whether by an explicit call to
+munmap() or via an internal unmap from exit() or exec() processing, we must
+munlock the pages if we're removing the last VM_LOCKED VMA that maps the pages.
+Before the unevictable/mlock changes, mlocking did not mark the pages in any
+way, so unmapping them required no processing.
+
+To munlock a range of memory under the unevictable/mlock infrastructure, the
+munmap() handler and task address space call tear down function
+munlock_vma_pages_all(). The name reflects the observation that one always
+specifies the entire VMA range when munlock()ing during unmap of a region.
+Because of the VMA filtering when mlocking() regions, only "normal" VMAs that
+actually contain mlocked pages will be passed to munlock_vma_pages_all().
+
+munlock_vma_pages_all() clears the VM_LOCKED VMA flag and, like mlock_fixup()
+for the munlock case, calls __munlock_vma_pages_range() to walk the page table
+for the VMA's memory range and munlock_vma_page() each resident page mapped by
+the VMA. This effectively munlocks the page, only if this is the last
+VM_LOCKED VMA that maps the page.
+
+
+try_to_unmap()
+--------------
+
+Pages can, of course, be mapped into multiple VMAs. Some of these VMAs may
+have VM_LOCKED flag set. It is possible for a page mapped into one or more
+VM_LOCKED VMAs not to have the PG_mlocked flag set and therefore reside on one
+of the active or inactive LRU lists. This could happen if, for example, a task
+in the process of munlocking the page could not isolate the page from the LRU.
+As a result, vmscan/shrink_page_list() might encounter such a page as described
+in section "vmscan's handling of unevictable pages". To handle this situation,
+try_to_unmap() checks for VM_LOCKED VMAs while it is walking a page's reverse
+map.
+
+try_to_unmap() is always called, by either vmscan for reclaim or for page
+migration, with the argument page locked and isolated from the LRU. Separate
+functions handle anonymous and mapped file and KSM pages, as these types of
+pages have different reverse map lookup mechanisms, with different locking.
+In each case, whether rmap_walk_anon() or rmap_walk_file() or rmap_walk_ksm(),
+it will call try_to_unmap_one() for every VMA which might contain the page.
+
+When trying to reclaim, if try_to_unmap_one() finds the page in a VM_LOCKED
+VMA, it will then mlock the page via mlock_vma_page() instead of unmapping it,
+and return SWAP_MLOCK to indicate that the page is unevictable: and the scan
+stops there.
+
+mlock_vma_page() is called while holding the page table's lock (in addition
+to the page lock, and the rmap lock): to serialize against concurrent mlock or
+munlock or munmap system calls, mm teardown (munlock_vma_pages_all), reclaim,
+holepunching, and truncation of file pages and their anonymous COWed pages.
+
+
+try_to_munlock() Reverse Map Scan
+---------------------------------
+
+.. warning::
+ [!] TODO/FIXME: a better name might be page_mlocked() - analogous to the
+ page_referenced() reverse map walker.
+
+When munlock_vma_page() [see section :ref:`munlock()/munlockall() System Call
+Handling <munlock_munlockall_handling>` above] tries to munlock a
+page, it needs to determine whether or not the page is mapped by any
+VM_LOCKED VMA without actually attempting to unmap all PTEs from the
+page. For this purpose, the unevictable/mlock infrastructure
+introduced a variant of try_to_unmap() called try_to_munlock().
+
+try_to_munlock() calls the same functions as try_to_unmap() for anonymous and
+mapped file and KSM pages with a flag argument specifying unlock versus unmap
+processing. Again, these functions walk the respective reverse maps looking
+for VM_LOCKED VMAs. When such a VMA is found, as in the try_to_unmap() case,
+the functions mlock the page via mlock_vma_page() and return SWAP_MLOCK. This
+undoes the pre-clearing of the page's PG_mlocked done by munlock_vma_page.
+
+Note that try_to_munlock()'s reverse map walk must visit every VMA in a page's
+reverse map to determine that a page is NOT mapped into any VM_LOCKED VMA.
+However, the scan can terminate when it encounters a VM_LOCKED VMA.
+Although try_to_munlock() might be called a great many times when munlocking a
+large region or tearing down a large address space that has been mlocked via
+mlockall(), overall this is a fairly rare event.
+
+
+Page Reclaim in shrink_*_list()
+-------------------------------
+
+shrink_active_list() culls any obviously unevictable pages - i.e.
+!page_evictable(page) - diverting these to the unevictable list.
+However, shrink_active_list() only sees unevictable pages that made it onto the
+active/inactive lru lists. Note that these pages do not have PageUnevictable
+set - otherwise they would be on the unevictable list and shrink_active_list
+would never see them.
+
+Some examples of these unevictable pages on the LRU lists are:
+
+ (1) ramfs pages that have been placed on the LRU lists when first allocated.
+
+ (2) SHM_LOCK'd shared memory pages. shmctl(SHM_LOCK) does not attempt to
+ allocate or fault in the pages in the shared memory region. This happens
+ when an application accesses the page the first time after SHM_LOCK'ing
+ the segment.
+
+ (3) mlocked pages that could not be isolated from the LRU and moved to the
+ unevictable list in mlock_vma_page().
+
+shrink_inactive_list() also diverts any unevictable pages that it finds on the
+inactive lists to the appropriate zone's unevictable list.
+
+shrink_inactive_list() should only see SHM_LOCK'd pages that became SHM_LOCK'd
+after shrink_active_list() had moved them to the inactive list, or pages mapped
+into VM_LOCKED VMAs that munlock_vma_page() couldn't isolate from the LRU to
+recheck via try_to_munlock(). shrink_inactive_list() won't notice the latter,
+but will pass on to shrink_page_list().
+
+shrink_page_list() again culls obviously unevictable pages that it could
+encounter for similar reason to shrink_inactive_list(). Pages mapped into
+VM_LOCKED VMAs but without PG_mlocked set will make it all the way to
+try_to_unmap(). shrink_page_list() will divert them to the unevictable list
+when try_to_unmap() returns SWAP_MLOCK, as discussed above.
diff --git a/Documentation/vm/z3fold.rst b/Documentation/vm/z3fold.rst
new file mode 100644
index 000000000..224e3c61d
--- /dev/null
+++ b/Documentation/vm/z3fold.rst
@@ -0,0 +1,30 @@
+.. _z3fold:
+
+======
+z3fold
+======
+
+z3fold is a special purpose allocator for storing compressed pages.
+It is designed to store up to three compressed pages per physical page.
+It is a zbud derivative which allows for higher compression
+ratio keeping the simplicity and determinism of its predecessor.
+
+The main differences between z3fold and zbud are:
+
+* unlike zbud, z3fold allows for up to PAGE_SIZE allocations
+* z3fold can hold up to 3 compressed pages in its page
+* z3fold doesn't export any API itself and is thus intended to be used
+ via the zpool API.
+
+To keep the determinism and simplicity, z3fold, just like zbud, always
+stores an integral number of compressed pages per page, but it can store
+up to 3 pages unlike zbud which can store at most 2. Therefore the
+compression ratio goes to around 2.7x while zbud's one is around 1.7x.
+
+Unlike zbud (but like zsmalloc for that matter) z3fold_alloc() does not
+return a dereferenceable pointer. Instead, it returns an unsigned long
+handle which encodes actual location of the allocated object.
+
+Keeping effective compression ratio close to zsmalloc's, z3fold doesn't
+depend on MMU enabled and provides more predictable reclaim behavior
+which makes it a better fit for small and response-critical systems.
diff --git a/Documentation/vm/zsmalloc.rst b/Documentation/vm/zsmalloc.rst
new file mode 100644
index 000000000..6e79893d6
--- /dev/null
+++ b/Documentation/vm/zsmalloc.rst
@@ -0,0 +1,82 @@
+.. _zsmalloc:
+
+========
+zsmalloc
+========
+
+This allocator is designed for use with zram. Thus, the allocator is
+supposed to work well under low memory conditions. In particular, it
+never attempts higher order page allocation which is very likely to
+fail under memory pressure. On the other hand, if we just use single
+(0-order) pages, it would suffer from very high fragmentation --
+any object of size PAGE_SIZE/2 or larger would occupy an entire page.
+This was one of the major issues with its predecessor (xvmalloc).
+
+To overcome these issues, zsmalloc allocates a bunch of 0-order pages
+and links them together using various 'struct page' fields. These linked
+pages act as a single higher-order page i.e. an object can span 0-order
+page boundaries. The code refers to these linked pages as a single entity
+called zspage.
+
+For simplicity, zsmalloc can only allocate objects of size up to PAGE_SIZE
+since this satisfies the requirements of all its current users (in the
+worst case, page is incompressible and is thus stored "as-is" i.e. in
+uncompressed form). For allocation requests larger than this size, failure
+is returned (see zs_malloc).
+
+Additionally, zs_malloc() does not return a dereferenceable pointer.
+Instead, it returns an opaque handle (unsigned long) which encodes actual
+location of the allocated object. The reason for this indirection is that
+zsmalloc does not keep zspages permanently mapped since that would cause
+issues on 32-bit systems where the VA region for kernel space mappings
+is very small. So, before using the allocating memory, the object has to
+be mapped using zs_map_object() to get a usable pointer and subsequently
+unmapped using zs_unmap_object().
+
+stat
+====
+
+With CONFIG_ZSMALLOC_STAT, we could see zsmalloc internal information via
+``/sys/kernel/debug/zsmalloc/<user name>``. Here is a sample of stat output::
+
+ # cat /sys/kernel/debug/zsmalloc/zram0/classes
+
+ class size almost_full almost_empty obj_allocated obj_used pages_used pages_per_zspage
+ ...
+ ...
+ 9 176 0 1 186 129 8 4
+ 10 192 1 0 2880 2872 135 3
+ 11 208 0 1 819 795 42 2
+ 12 224 0 1 219 159 12 4
+ ...
+ ...
+
+
+class
+ index
+size
+ object size zspage stores
+almost_empty
+ the number of ZS_ALMOST_EMPTY zspages(see below)
+almost_full
+ the number of ZS_ALMOST_FULL zspages(see below)
+obj_allocated
+ the number of objects allocated
+obj_used
+ the number of objects allocated to the user
+pages_used
+ the number of pages allocated for the class
+pages_per_zspage
+ the number of 0-order pages to make a zspage
+
+We assign a zspage to ZS_ALMOST_EMPTY fullness group when n <= N / f, where
+
+* n = number of allocated objects
+* N = total number of objects zspage can store
+* f = fullness_threshold_frac(ie, 4 at the moment)
+
+Similarly, we assign zspage to:
+
+* ZS_ALMOST_FULL when n > N / f
+* ZS_EMPTY when n == 0
+* ZS_FULL when n == N
diff --git a/Documentation/vm/zswap.rst b/Documentation/vm/zswap.rst
new file mode 100644
index 000000000..d8d9fa4a1
--- /dev/null
+++ b/Documentation/vm/zswap.rst
@@ -0,0 +1,152 @@
+.. _zswap:
+
+=====
+zswap
+=====
+
+Overview
+========
+
+Zswap is a lightweight compressed cache for swap pages. It takes pages that are
+in the process of being swapped out and attempts to compress them into a
+dynamically allocated RAM-based memory pool. zswap basically trades CPU cycles
+for potentially reduced swap I/O.  This trade-off can also result in a
+significant performance improvement if reads from the compressed cache are
+faster than reads from a swap device.
+
+.. note::
+ Zswap is a new feature as of v3.11 and interacts heavily with memory
+ reclaim. This interaction has not been fully explored on the large set of
+ potential configurations and workloads that exist. For this reason, zswap
+ is a work in progress and should be considered experimental.
+
+ Some potential benefits:
+
+* Desktop/laptop users with limited RAM capacities can mitigate the
+ performance impact of swapping.
+* Overcommitted guests that share a common I/O resource can
+ dramatically reduce their swap I/O pressure, avoiding heavy handed I/O
+ throttling by the hypervisor. This allows more work to get done with less
+ impact to the guest workload and guests sharing the I/O subsystem
+* Users with SSDs as swap devices can extend the life of the device by
+ drastically reducing life-shortening writes.
+
+Zswap evicts pages from compressed cache on an LRU basis to the backing swap
+device when the compressed pool reaches its size limit. This requirement had
+been identified in prior community discussions.
+
+Whether Zswap is enabled at the boot time depends on whether
+the ``CONFIG_ZSWAP_DEFAULT_ON`` Kconfig option is enabled or not.
+This setting can then be overridden by providing the kernel command line
+``zswap.enabled=`` option, for example ``zswap.enabled=0``.
+Zswap can also be enabled and disabled at runtime using the sysfs interface.
+An example command to enable zswap at runtime, assuming sysfs is mounted
+at ``/sys``, is::
+
+ echo 1 > /sys/module/zswap/parameters/enabled
+
+When zswap is disabled at runtime it will stop storing pages that are
+being swapped out. However, it will _not_ immediately write out or fault
+back into memory all of the pages stored in the compressed pool. The
+pages stored in zswap will remain in the compressed pool until they are
+either invalidated or faulted back into memory. In order to force all
+pages out of the compressed pool, a swapoff on the swap device(s) will
+fault back into memory all swapped out pages, including those in the
+compressed pool.
+
+Design
+======
+
+Zswap receives pages for compression through the Frontswap API and is able to
+evict pages from its own compressed pool on an LRU basis and write them back to
+the backing swap device in the case that the compressed pool is full.
+
+Zswap makes use of zpool for the managing the compressed memory pool. Each
+allocation in zpool is not directly accessible by address. Rather, a handle is
+returned by the allocation routine and that handle must be mapped before being
+accessed. The compressed memory pool grows on demand and shrinks as compressed
+pages are freed. The pool is not preallocated. By default, a zpool
+of type selected in ``CONFIG_ZSWAP_ZPOOL_DEFAULT`` Kconfig option is created,
+but it can be overridden at boot time by setting the ``zpool`` attribute,
+e.g. ``zswap.zpool=zbud``. It can also be changed at runtime using the sysfs
+``zpool`` attribute, e.g.::
+
+ echo zbud > /sys/module/zswap/parameters/zpool
+
+The zbud type zpool allocates exactly 1 page to store 2 compressed pages, which
+means the compression ratio will always be 2:1 or worse (because of half-full
+zbud pages). The zsmalloc type zpool has a more complex compressed page
+storage method, and it can achieve greater storage densities. However,
+zsmalloc does not implement compressed page eviction, so once zswap fills it
+cannot evict the oldest page, it can only reject new pages.
+
+When a swap page is passed from frontswap to zswap, zswap maintains a mapping
+of the swap entry, a combination of the swap type and swap offset, to the zpool
+handle that references that compressed swap page. This mapping is achieved
+with a red-black tree per swap type. The swap offset is the search key for the
+tree nodes.
+
+During a page fault on a PTE that is a swap entry, frontswap calls the zswap
+load function to decompress the page into the page allocated by the page fault
+handler.
+
+Once there are no PTEs referencing a swap page stored in zswap (i.e. the count
+in the swap_map goes to 0) the swap code calls the zswap invalidate function,
+via frontswap, to free the compressed entry.
+
+Zswap seeks to be simple in its policies. Sysfs attributes allow for one user
+controlled policy:
+
+* max_pool_percent - The maximum percentage of memory that the compressed
+ pool can occupy.
+
+The default compressor is selected in ``CONFIG_ZSWAP_COMPRESSOR_DEFAULT``
+Kconfig option, but it can be overridden at boot time by setting the
+``compressor`` attribute, e.g. ``zswap.compressor=lzo``.
+It can also be changed at runtime using the sysfs "compressor"
+attribute, e.g.::
+
+ echo lzo > /sys/module/zswap/parameters/compressor
+
+When the zpool and/or compressor parameter is changed at runtime, any existing
+compressed pages are not modified; they are left in their own zpool. When a
+request is made for a page in an old zpool, it is uncompressed using its
+original compressor. Once all pages are removed from an old zpool, the zpool
+and its compressor are freed.
+
+Some of the pages in zswap are same-value filled pages (i.e. contents of the
+page have same value or repetitive pattern). These pages include zero-filled
+pages and they are handled differently. During store operation, a page is
+checked if it is a same-value filled page before compressing it. If true, the
+compressed length of the page is set to zero and the pattern or same-filled
+value is stored.
+
+Same-value filled pages identification feature is enabled by default and can be
+disabled at boot time by setting the ``same_filled_pages_enabled`` attribute
+to 0, e.g. ``zswap.same_filled_pages_enabled=0``. It can also be enabled and
+disabled at runtime using the sysfs ``same_filled_pages_enabled``
+attribute, e.g.::
+
+ echo 1 > /sys/module/zswap/parameters/same_filled_pages_enabled
+
+When zswap same-filled page identification is disabled at runtime, it will stop
+checking for the same-value filled pages during store operation. However, the
+existing pages which are marked as same-value filled pages remain stored
+unchanged in zswap until they are either loaded or invalidated.
+
+To prevent zswap from shrinking pool when zswap is full and there's a high
+pressure on swap (this will result in flipping pages in and out zswap pool
+without any real benefit but with a performance drop for the system), a
+special parameter has been introduced to implement a sort of hysteresis to
+refuse taking pages into zswap pool until it has sufficient space if the limit
+has been hit. To set the threshold at which zswap would start accepting pages
+again after it became full, use the sysfs ``accept_threshold_percent``
+attribute, e. g.::
+
+ echo 80 > /sys/module/zswap/parameters/accept_threshold_percent
+
+Setting this parameter to 100 will disable the hysteresis.
+
+A debugfs interface is provided for various statistic about pool size, number
+of pages stored, same-value filled pages and various counters for the reasons
+pages are rejected.