Adding upstream version 6.1.76.upstream/6.1.76

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

45 files changed, 5698 insertions, 0 deletions

diff --git a/Documentation/mm/active_mm.rst b/Documentation/mm/active_mm.rst
new file mode 100644
index 000000000..6f8269c28
--- /dev/null
+++ b/Documentation/mm/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/mm/arch_pgtable_helpers.rst b/Documentation/mm/arch_pgtable_helpers.rst
new file mode 100644
index 000000000..cbaee9e59
--- /dev/null
+++ b/Documentation/mm/arch_pgtable_helpers.rst
@@ -0,0 +1,260 @@
+.. 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_clear                | Clears a PTE                                     |
++---------------------------+--------------------------------------------------+
+| ptep_get_and_clear        | Clears and returns PTE                           |
++---------------------------+--------------------------------------------------+
+| ptep_get_and_clear_full   | Clears and returns PTE (batched PTE unmap)       |
++---------------------------+--------------------------------------------------+
+| 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_writable_migration_entry   | Tests a write migration swapped entry        |
++-------------------------------+----------------------------------------------+
+| make_readable_migration_entry | Creates a read migration swapped entry       |
++-------------------------------+----------------------------------------------+
+| make_writable_migration_entry | Creates a write migration swapped entry      |
++-------------------------------+----------------------------------------------+
+
+[1] https://lore.kernel.org/linux-mm/20181017020930.GN30832@redhat.com/
diff --git a/Documentation/mm/balance.rst b/Documentation/mm/balance.rst
new file mode 100644
index 000000000..6a1fadf3e
--- /dev/null
+++ b/Documentation/mm/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/mm/bootmem.rst b/Documentation/mm/bootmem.rst
new file mode 100644
index 000000000..eb2b31eed
--- /dev/null
+++ b/Documentation/mm/bootmem.rst
@@ -0,0 +1,5 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+===========
+Boot Memory
+===========
diff --git a/Documentation/mm/damon/api.rst b/Documentation/mm/damon/api.rst
new file mode 100644
index 000000000..08f34df45
--- /dev/null
+++ b/Documentation/mm/damon/api.rst
@@ -0,0 +1,20 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=============
+API Reference
+=============
+
+Kernel space programs can use every feature of DAMON using below APIs.  All you
+need to do is including ``damon.h``, which is located in ``include/linux/`` of
+the source tree.
+
+Structures
+==========
+
+.. kernel-doc:: include/linux/damon.h
+
+
+Functions
+=========
+
+.. kernel-doc:: mm/damon/core.c
diff --git a/Documentation/mm/damon/design.rst b/Documentation/mm/damon/design.rst
new file mode 100644
index 000000000..0cff6fac6
--- /dev/null
+++ b/Documentation/mm/damon/design.rst
@@ -0,0 +1,176 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+======
+Design
+======
+
+Configurable Layers
+===================
+
+DAMON provides data access monitoring functionality while making the accuracy
+and the overhead controllable.  The fundamental access monitorings require
+primitives that dependent on and optimized for the target address space.  On
+the other hand, the accuracy and overhead tradeoff mechanism, which is the core
+of DAMON, is in the pure logic space.  DAMON separates the two parts in
+different layers and defines its interface to allow various low level
+primitives implementations configurable with the core logic.  We call the low
+level primitives implementations monitoring operations.
+
+Due to this separated design and the configurable interface, users can extend
+DAMON for any address space by configuring the core logics with appropriate
+monitoring operations.  If appropriate one is not provided, users can implement
+the operations on their own.
+
+For example, physical memory, virtual memory, swap space, those for specific
+processes, NUMA nodes, files, and backing memory devices would be supportable.
+Also, if some architectures or devices support special optimized access check
+primitives, those will be easily configurable.
+
+
+Reference Implementations of Address Space Specific Monitoring Operations
+=========================================================================
+
+The monitoring operations are defined in two parts:
+
+1. Identification of the monitoring target address range for the address space.
+2. Access check of specific address range in the target space.
+
+DAMON currently provides the implementations of the operations for the physical
+and virtual address spaces. Below two subsections describe how those work.
+
+
+VMA-based Target Address Range Construction
+-------------------------------------------
+
+This is only for the virtual address space monitoring operations
+implementation.  That for the physical address space simply asks users to
+manually set the monitoring target address ranges.
+
+Only small parts in the super-huge virtual address space of the processes are
+mapped to the physical memory and accessed.  Thus, tracking the unmapped
+address regions is just wasteful.  However, because DAMON can deal with some
+level of noise using the adaptive regions adjustment mechanism, tracking every
+mapping is not strictly required but could even incur a high overhead in some
+cases.  That said, too huge unmapped areas inside the monitoring target should
+be removed to not take the time for the adaptive mechanism.
+
+For the reason, this implementation converts the complex mappings to three
+distinct regions that cover every mapped area of the address space.  The two
+gaps between the three regions are the two biggest unmapped areas in the given
+address space.  The two biggest unmapped areas would be the gap between the
+heap and the uppermost mmap()-ed region, and the gap between the lowermost
+mmap()-ed region and the stack in most of the cases.  Because these gaps are
+exceptionally huge in usual address spaces, excluding these will be sufficient
+to make a reasonable trade-off.  Below shows this in detail::
+
+    <heap>
+    <BIG UNMAPPED REGION 1>
+    <uppermost mmap()-ed region>
+    (small mmap()-ed regions and munmap()-ed regions)
+    <lowermost mmap()-ed region>
+    <BIG UNMAPPED REGION 2>
+    <stack>
+
+
+PTE Accessed-bit Based Access Check
+-----------------------------------
+
+Both of the implementations for physical and virtual address spaces use PTE
+Accessed-bit for basic access checks.  Only one difference is the way of
+finding the relevant PTE Accessed bit(s) from the address.  While the
+implementation for the virtual address walks the page table for the target task
+of the address, the implementation for the physical address walks every page
+table having a mapping to the address.  In this way, the implementations find
+and clear the bit(s) for next sampling target address and checks whether the
+bit(s) set again after one sampling period.  This could disturb other kernel
+subsystems using the Accessed bits, namely Idle page tracking and the reclaim
+logic.  DAMON does nothing to avoid disturbing Idle page tracking, so handling
+the interference is the responsibility of sysadmins.  However, it solves the
+conflict with the reclaim logic using ``PG_idle`` and ``PG_young`` page flags,
+as Idle page tracking does.
+
+
+Address Space Independent Core Mechanisms
+=========================================
+
+Below four sections describe each of the DAMON core mechanisms and the five
+monitoring attributes, ``sampling interval``, ``aggregation interval``,
+``update interval``, ``minimum number of regions``, and ``maximum number of
+regions``.
+
+
+Access Frequency Monitoring
+---------------------------
+
+The output of DAMON says what pages are how frequently accessed for a given
+duration.  The resolution of the access frequency is controlled by setting
+``sampling interval`` and ``aggregation interval``.  In detail, DAMON checks
+access to each page per ``sampling interval`` and aggregates the results.  In
+other words, counts the number of the accesses to each page.  After each
+``aggregation interval`` passes, DAMON calls callback functions that previously
+registered by users so that users can read the aggregated results and then
+clears the results.  This can be described in below simple pseudo-code::
+
+    while monitoring_on:
+        for page in monitoring_target:
+            if accessed(page):
+                nr_accesses[page] += 1
+        if time() % aggregation_interval == 0:
+            for callback in user_registered_callbacks:
+                callback(monitoring_target, nr_accesses)
+            for page in monitoring_target:
+                nr_accesses[page] = 0
+        sleep(sampling interval)
+
+The monitoring overhead of this mechanism will arbitrarily increase as the
+size of the target workload grows.
+
+
+Region Based Sampling
+---------------------
+
+To avoid the unbounded increase of the overhead, DAMON groups adjacent pages
+that assumed to have the same access frequencies into a region.  As long as the
+assumption (pages in a region have the same access frequencies) is kept, only
+one page in the region is required to be checked.  Thus, for each ``sampling
+interval``, DAMON randomly picks one page in each region, waits for one
+``sampling interval``, checks whether the page is accessed meanwhile, and
+increases the access frequency of the region if so.  Therefore, the monitoring
+overhead is controllable by setting the number of regions.  DAMON allows users
+to set the minimum and the maximum number of regions for the trade-off.
+
+This scheme, however, cannot preserve the quality of the output if the
+assumption is not guaranteed.
+
+
+Adaptive Regions Adjustment
+---------------------------
+
+Even somehow the initial monitoring target regions are well constructed to
+fulfill the assumption (pages in same region have similar access frequencies),
+the data access pattern can be dynamically changed.  This will result in low
+monitoring quality.  To keep the assumption as much as possible, DAMON
+adaptively merges and splits each region based on their access frequency.
+
+For each ``aggregation interval``, it compares the access frequencies of
+adjacent regions and merges those if the frequency difference is small.  Then,
+after it reports and clears the aggregated access frequency of each region, it
+splits each region into two or three regions if the total number of regions
+will not exceed the user-specified maximum number of regions after the split.
+
+In this way, DAMON provides its best-effort quality and minimal overhead while
+keeping the bounds users set for their trade-off.
+
+
+Dynamic Target Space Updates Handling
+-------------------------------------
+
+The monitoring target address range could dynamically changed.  For example,
+virtual memory could be dynamically mapped and unmapped.  Physical memory could
+be hot-plugged.
+
+As the changes could be quite frequent in some cases, DAMON allows the
+monitoring operations to check dynamic changes including memory mapping changes
+and applies it to monitoring operations-related data structures such as the
+abstracted monitoring target memory area only for each of a user-specified time
+interval (``update interval``).
diff --git a/Documentation/mm/damon/faq.rst b/Documentation/mm/damon/faq.rst
new file mode 100644
index 000000000..dde7e2414
--- /dev/null
+++ b/Documentation/mm/damon/faq.rst
@@ -0,0 +1,50 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+==========================
+Frequently Asked Questions
+==========================
+
+Why a new subsystem, instead of extending perf or other user space tools?
+=========================================================================
+
+First, because it needs to be lightweight as much as possible so that it can be
+used online, any unnecessary overhead such as kernel - user space context
+switching cost should be avoided.  Second, DAMON aims to be used by other
+programs including the kernel.  Therefore, having a dependency on specific
+tools like perf is not desirable.  These are the two biggest reasons why DAMON
+is implemented in the kernel space.
+
+
+Can 'idle pages tracking' or 'perf mem' substitute DAMON?
+=========================================================
+
+Idle page tracking is a low level primitive for access check of the physical
+address space.  'perf mem' is similar, though it can use sampling to minimize
+the overhead.  On the other hand, DAMON is a higher-level framework for the
+monitoring of various address spaces.  It is focused on memory management
+optimization and provides sophisticated accuracy/overhead handling mechanisms.
+Therefore, 'idle pages tracking' and 'perf mem' could provide a subset of
+DAMON's output, but cannot substitute DAMON.
+
+
+Does DAMON support virtual memory only?
+=======================================
+
+No.  The core of the DAMON is address space independent.  The address space
+specific monitoring operations including monitoring target regions
+constructions and actual access checks can be implemented and configured on the
+DAMON core by the users.  In this way, DAMON users can monitor any address
+space with any access check technique.
+
+Nonetheless, DAMON provides vma/rmap tracking and PTE Accessed bit check based
+implementations of the address space dependent functions for the virtual memory
+and the physical memory by default, for a reference and convenient use.
+
+
+Can I simply monitor page granularity?
+======================================
+
+Yes.  You can do so by setting the ``min_nr_regions`` attribute higher than the
+working set size divided by the page size.  Because the monitoring target
+regions size is forced to be ``>=page size``, the region split will make no
+effect.
diff --git a/Documentation/mm/damon/index.rst b/Documentation/mm/damon/index.rst
new file mode 100644
index 000000000..48c0bbff9
--- /dev/null
+++ b/Documentation/mm/damon/index.rst
@@ -0,0 +1,29 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+==========================
+DAMON: Data Access MONitor
+==========================
+
+DAMON is a data access monitoring framework subsystem for the Linux kernel.
+The core mechanisms of DAMON (refer to :doc:`design` for the detail) make it
+
+ - *accurate* (the monitoring output is useful enough for DRAM level memory
+   management; It might not appropriate for CPU Cache levels, though),
+ - *light-weight* (the monitoring overhead is low enough to be applied online),
+   and
+ - *scalable* (the upper-bound of the overhead is in constant range regardless
+   of the size of target workloads).
+
+Using this framework, therefore, the kernel's memory management mechanisms can
+make advanced decisions.  Experimental memory management optimization works
+that incurring high data accesses monitoring overhead could implemented again.
+In user space, meanwhile, users who have some special workloads can write
+personalized applications for better understanding and optimizations of their
+workloads and systems.
+
+.. toctree::
+   :maxdepth: 2
+
+   faq
+   design
+   api
diff --git a/Documentation/mm/free_page_reporting.rst b/Documentation/mm/free_page_reporting.rst
new file mode 100644
index 000000000..8c05e62d8
--- /dev/null
+++ b/Documentation/mm/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/mm/frontswap.rst b/Documentation/mm/frontswap.rst
new file mode 100644
index 000000000..feecc5e24
--- /dev/null
+++ b/Documentation/mm/frontswap.rst
@@ -0,0 +1,266 @@
+.. _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.
+
+.. _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.
+
+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".
+
+Frontswap 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.
+
+* 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/mm/highmem.rst b/Documentation/mm/highmem.rst
new file mode 100644
index 000000000..0f731d919
--- /dev/null
+++ b/Documentation/mm/highmem.rst
@@ -0,0 +1,190 @@
+.. _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. The following
+list shows them in order of preference of use.
+
+* kmap_local_page().  This function is used to require short term mappings.
+  It can be invoked from any context (including interrupts) but the mappings
+  can only be used in the context which acquired them.
+
+  This function should be preferred, where feasible, over all the others.
+
+  These mappings are thread-local and CPU-local, meaning that the mapping
+  can only be accessed from within this thread and the thread is bound to the
+  CPU while the mapping is active. Although preemption is never disabled by
+  this function, the CPU can not be unplugged from the system via
+  CPU-hotplug until the mapping is disposed.
+
+  It's valid to take pagefaults in a local kmap region, unless the context
+  in which the local mapping is acquired does not allow it for other reasons.
+
+  As said, pagefaults and preemption are never disabled. There is no need to
+  disable preemption because, when context switches to a different task, the
+  maps of the outgoing task are saved and those of the incoming one are
+  restored.
+
+  kmap_local_page() always returns a valid virtual address and it is assumed
+  that kunmap_local() will never fail.
+
+  On CONFIG_HIGHMEM=n kernels and for low memory pages this returns the
+  virtual address of the direct mapping. Only real highmem pages are
+  temporarily mapped. Therefore, users may call a plain page_address()
+  for pages which are known to not come from ZONE_HIGHMEM. However, it is
+  always safe to use kmap_local_page() / kunmap_local().
+
+  While it is significantly faster than kmap(), for the higmem case it
+  comes with restrictions about the pointers validity. Contrary to kmap()
+  mappings, the local mappings are only valid in the context of the caller
+  and cannot be handed to other contexts. This implies that users must
+  be absolutely sure to keep the use of the return address local to the
+  thread which mapped it.
+
+  Most code can be designed to use thread local mappings. User should
+  therefore try to design their code to avoid the use of kmap() by mapping
+  pages in the same thread the address will be used and prefer
+  kmap_local_page().
+
+  Nesting kmap_local_page() and kmap_atomic() mappings is allowed to a certain
+  extent (up to KMAP_TYPE_NR) but their invocations have to be strictly ordered
+  because the map implementation is stack based. See kmap_local_page() kdocs
+  (included in the "Functions" section) for details on how to manage nested
+  mappings.
+
+* 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 does not
+  sleep and the callers too may not sleep until after kunmap_atomic() is
+  called.
+
+  Each call of kmap_atomic() in the kernel creates a non-preemptible section
+  and disable pagefaults. This could be a source of unwanted latency. Therefore
+  users should prefer kmap_local_page() instead of kmap_atomic().
+
+  It is assumed that k[un]map_atomic() won't fail.
+
+* kmap().  This should be used to make short duration mapping of a single
+  page with no restrictions on preemption or migration. It comes with an
+  overhead as mapping space is restricted and protected by a global lock
+  for synchronization. When mapping is no longer needed, the address that
+  the page was mapped to must be released with kunmap().
+
+  Mapping changes must be propagated across all the CPUs. kmap() also
+  requires global TLB invalidation when the kmap's pool wraps and it might
+  block when the mapping space is fully utilized until a slot becomes
+  available. Therefore, kmap() is only callable from preemptible context.
+
+  All the above work is necessary if a mapping must last for a relatively
+  long time but the bulk of high-memory mappings in the kernel are
+  short-lived and only used in one place. This means that the cost of
+  kmap() is mostly wasted in such cases. kmap() was not intended for long
+  term mappings but it has morphed in that direction and its use is
+  strongly discouraged in newer code and the set of the preceding functions
+  should be preferred.
+
+  On 64-bit systems, calls to kmap_local_page(), kmap_atomic() and kmap() have
+  no real work to do because a 64-bit address space is more than sufficient to
+  address all the physical memory whose pages are permanently mapped.
+
+* 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.
+
+
+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.
+
+
+Functions
+=========
+
+.. kernel-doc:: include/linux/highmem.h
+.. kernel-doc:: include/linux/highmem-internal.h
diff --git a/Documentation/mm/hmm.rst b/Documentation/mm/hmm.rst
new file mode 100644
index 000000000..f2a59ed82
--- /dev/null
+++ b/Documentation/mm/hmm.rst
@@ -0,0 +1,452 @@
+.. _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 ``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));
+
+   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.
+
+Exclusive access memory
+=======================
+
+Some devices have features such as atomic PTE bits that can be used to implement
+atomic access to system memory. To support atomic operations to a shared virtual
+memory page such a device needs access to that page which is exclusive of any
+userspace access from the CPU. The ``make_device_exclusive_range()`` function
+can be used to make a memory range inaccessible from userspace.
+
+This replaces all mappings for pages in the given range with special swap
+entries. Any attempt to access the swap entry results in a fault which is
+resovled by replacing the entry with the original mapping. A driver gets
+notified that the mapping has been changed by MMU notifiers, after which point
+it will no longer have exclusive access to the page. Exclusive access is
+guranteed to last until the driver drops the page lock and page reference, at
+which point any CPU faults on the page may proceed as described.
+
+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/mm/hugetlbfs_reserv.rst b/Documentation/mm/hugetlbfs_reserv.rst
new file mode 100644
index 000000000..f143954e0
--- /dev/null
+++ b/Documentation/mm/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/mm/hwpoison.rst b/Documentation/mm/hwpoison.rst
new file mode 100644
index 000000000..b9d5253c1
--- /dev/null
+++ b/Documentation/mm/hwpoison.rst
@@ -0,0 +1,184 @@
+.. 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.
+
+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. Once any hardware
+	memory failure happens, this feature is disabled.
+
+  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.
+
+---
+Andi Kleen, Oct 2009
diff --git a/Documentation/mm/index.rst b/Documentation/mm/index.rst
new file mode 100644
index 000000000..4aa12b8be
--- /dev/null
+++ b/Documentation/mm/index.rst
@@ -0,0 +1,69 @@
+=====================================
+Linux Memory Management Documentation
+=====================================
+
+Memory Management Guide
+=======================
+
+This is a guide to understanding the memory management subsystem
+of Linux.  If you are looking for advice on simply allocating memory,
+see the :ref:`memory_allocation`.  For controlling and tuning guides,
+see the :doc:`admin guide <../admin-guide/mm/index>`.
+
+.. toctree::
+   :maxdepth: 1
+
+   physical_memory
+   page_tables
+   process_addrs
+   bootmem
+   page_allocation
+   vmalloc
+   slab
+   highmem
+   page_reclaim
+   swap
+   page_cache
+   shmfs
+   oom
+
+Legacy Documentation
+====================
+
+This is a collection of older documents about the Linux memory management
+(MM) subsystem internals with different level of details ranging from
+notes and mailing list responses for elaborating descriptions of data
+structures and algorithms.  It should all be integrated nicely into the
+above structured documentation, or deleted if it has served its purpose.
+
+.. toctree::
+   :maxdepth: 1
+
+   active_mm
+   arch_pgtable_helpers
+   balance
+   damon/index
+   free_page_reporting
+   frontswap
+   hmm
+   hwpoison
+   hugetlbfs_reserv
+   ksm
+   memory-model
+   mmu_notifier
+   multigen_lru
+   numa
+   overcommit-accounting
+   page_migration
+   page_frags
+   page_owner
+   page_table_check
+   remap_file_pages
+   slub
+   split_page_table_lock
+   transhuge
+   unevictable-lru
+   vmalloced-kernel-stacks
+   vmemmap_dedup
+   z3fold
+   zsmalloc
diff --git a/Documentation/mm/ksm.rst b/Documentation/mm/ksm.rst
new file mode 100644
index 000000000..f83cfbc12
--- /dev/null
+++ b/Documentation/mm/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 ksm_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/mm/memory-model.rst b/Documentation/mm/memory-model.rst
new file mode 100644
index 000000000..3779e562d
--- /dev/null
+++ b/Documentation/mm/memory-model.rst
@@ -0,0 +1,177 @@
+.. 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 two memory models:
+FLATMEM 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.
+
+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.
+
+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/mm/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/mm/mmu_notifier.rst b/Documentation/mm/mmu_notifier.rst
new file mode 100644
index 000000000..df5d7777f
--- /dev/null
+++ b/Documentation/mm/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/mm/multigen_lru.rst b/Documentation/mm/multigen_lru.rst
new file mode 100644
index 000000000..d8f721f98
--- /dev/null
+++ b/Documentation/mm/multigen_lru.rst
@@ -0,0 +1,159 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=============
+Multi-Gen LRU
+=============
+The multi-gen LRU is an alternative LRU implementation that optimizes
+page reclaim and improves performance under memory pressure. Page
+reclaim decides the kernel's caching policy and ability to overcommit
+memory. It directly impacts the kswapd CPU usage and RAM efficiency.
+
+Design overview
+===============
+Objectives
+----------
+The design objectives are:
+
+* Good representation of access recency
+* Try to profit from spatial locality
+* Fast paths to make obvious choices
+* Simple self-correcting heuristics
+
+The representation of access recency is at the core of all LRU
+implementations. In the multi-gen LRU, each generation represents a
+group of pages with similar access recency. Generations establish a
+(time-based) common frame of reference and therefore help make better
+choices, e.g., between different memcgs on a computer or different
+computers in a data center (for job scheduling).
+
+Exploiting spatial locality improves efficiency when gathering the
+accessed bit. A rmap walk targets a single page and does not try to
+profit from discovering a young PTE. A page table walk can sweep all
+the young PTEs in an address space, but the address space can be too
+sparse to make a profit. The key is to optimize both methods and use
+them in combination.
+
+Fast paths reduce code complexity and runtime overhead. Unmapped pages
+do not require TLB flushes; clean pages do not require writeback.
+These facts are only helpful when other conditions, e.g., access
+recency, are similar. With generations as a common frame of reference,
+additional factors stand out. But obvious choices might not be good
+choices; thus self-correction is necessary.
+
+The benefits of simple self-correcting heuristics are self-evident.
+Again, with generations as a common frame of reference, this becomes
+attainable. Specifically, pages in the same generation can be
+categorized based on additional factors, and a feedback loop can
+statistically compare the refault percentages across those categories
+and infer which of them are better choices.
+
+Assumptions
+-----------
+The protection of hot pages and the selection of cold pages are based
+on page access channels and patterns. There are two access channels:
+
+* Accesses through page tables
+* Accesses through file descriptors
+
+The protection of the former channel is by design stronger because:
+
+1. The uncertainty in determining the access patterns of the former
+   channel is higher due to the approximation of the accessed bit.
+2. The cost of evicting the former channel is higher due to the TLB
+   flushes required and the likelihood of encountering the dirty bit.
+3. The penalty of underprotecting the former channel is higher because
+   applications usually do not prepare themselves for major page
+   faults like they do for blocked I/O. E.g., GUI applications
+   commonly use dedicated I/O threads to avoid blocking rendering
+   threads.
+
+There are also two access patterns:
+
+* Accesses exhibiting temporal locality
+* Accesses not exhibiting temporal locality
+
+For the reasons listed above, the former channel is assumed to follow
+the former pattern unless ``VM_SEQ_READ`` or ``VM_RAND_READ`` is
+present, and the latter channel is assumed to follow the latter
+pattern unless outlying refaults have been observed.
+
+Workflow overview
+=================
+Evictable pages are divided into multiple generations for each
+``lruvec``. The youngest generation number is stored in
+``lrugen->max_seq`` for both anon and file types as they are aged on
+an equal footing. The oldest generation numbers are stored in
+``lrugen->min_seq[]`` separately for anon and file types as clean file
+pages can be evicted regardless of swap constraints. These three
+variables are monotonically increasing.
+
+Generation numbers are truncated into ``order_base_2(MAX_NR_GENS+1)``
+bits in order to fit into the gen counter in ``folio->flags``. Each
+truncated generation number is an index to ``lrugen->folios[]``. The
+sliding window technique is used to track at least ``MIN_NR_GENS`` and
+at most ``MAX_NR_GENS`` generations. The gen counter stores a value
+within ``[1, MAX_NR_GENS]`` while a page is on one of
+``lrugen->folios[]``; otherwise it stores zero.
+
+Each generation is divided into multiple tiers. A page accessed ``N``
+times through file descriptors is in tier ``order_base_2(N)``. Unlike
+generations, tiers do not have dedicated ``lrugen->folios[]``. In
+contrast to moving across generations, which requires the LRU lock,
+moving across tiers only involves atomic operations on
+``folio->flags`` and therefore has a negligible cost. A feedback loop
+modeled after the PID controller monitors refaults over all the tiers
+from anon and file types and decides which tiers from which types to
+evict or protect.
+
+There are two conceptually independent procedures: the aging and the
+eviction. They form a closed-loop system, i.e., the page reclaim.
+
+Aging
+-----
+The aging produces young generations. Given an ``lruvec``, it
+increments ``max_seq`` when ``max_seq-min_seq+1`` approaches
+``MIN_NR_GENS``. The aging promotes hot pages to the youngest
+generation when it finds them accessed through page tables; the
+demotion of cold pages happens consequently when it increments
+``max_seq``. The aging uses page table walks and rmap walks to find
+young PTEs. For the former, it iterates ``lruvec_memcg()->mm_list``
+and calls ``walk_page_range()`` with each ``mm_struct`` on this list
+to scan PTEs, and after each iteration, it increments ``max_seq``. For
+the latter, when the eviction walks the rmap and finds a young PTE,
+the aging scans the adjacent PTEs. For both, on finding a young PTE,
+the aging clears the accessed bit and updates the gen counter of the
+page mapped by this PTE to ``(max_seq%MAX_NR_GENS)+1``.
+
+Eviction
+--------
+The eviction consumes old generations. Given an ``lruvec``, it
+increments ``min_seq`` when ``lrugen->folios[]`` indexed by
+``min_seq%MAX_NR_GENS`` becomes empty. To select a type and a tier to
+evict from, it first compares ``min_seq[]`` to select the older type.
+If both types are equally old, it selects the one whose first tier has
+a lower refault percentage. The first tier contains single-use
+unmapped clean pages, which are the best bet. The eviction sorts a
+page according to its gen counter if the aging has found this page
+accessed through page tables and updated its gen counter. It also
+moves a page to the next generation, i.e., ``min_seq+1``, if this page
+was accessed multiple times through file descriptors and the feedback
+loop has detected outlying refaults from the tier this page is in. To
+this end, the feedback loop uses the first tier as the baseline, for
+the reason stated earlier.
+
+Summary
+-------
+The multi-gen LRU can be disassembled into the following parts:
+
+* Generations
+* Rmap walks
+* Page table walks
+* Bloom filters
+* PID controller
+
+The aging and the eviction form a producer-consumer model;
+specifically, the latter drives the former by the sliding window over
+generations. Within the aging, rmap walks drive page table walks by
+inserting hot densely populated page tables to the Bloom filters.
+Within the eviction, the PID controller uses refaults as the feedback
+to select types to evict and tiers to protect.
diff --git a/Documentation/mm/numa.rst b/Documentation/mm/numa.rst
new file mode 100644
index 000000000..99fdeca91
--- /dev/null
+++ b/Documentation/mm/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/mm/oom.rst b/Documentation/mm/oom.rst
new file mode 100644
index 000000000..18e9e40c1
--- /dev/null
+++ b/Documentation/mm/oom.rst
@@ -0,0 +1,5 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+======================
+Out Of Memory Handling
+======================
diff --git a/Documentation/mm/overcommit-accounting.rst b/Documentation/mm/overcommit-accounting.rst
new file mode 100644
index 000000000..a4895d6fc
--- /dev/null
+++ b/Documentation/mm/overcommit-accounting.rst
@@ -0,0 +1,86 @@
+=====================
+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). These only have an effect
+when ``vm.overcommit_memory`` is set to 2.
+
+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/mm/page_allocation.rst b/Documentation/mm/page_allocation.rst
new file mode 100644
index 000000000..d9b449556
--- /dev/null
+++ b/Documentation/mm/page_allocation.rst
@@ -0,0 +1,5 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+===============
+Page Allocation
+===============
diff --git a/Documentation/mm/page_cache.rst b/Documentation/mm/page_cache.rst
new file mode 100644
index 000000000..75eba7c43
--- /dev/null
+++ b/Documentation/mm/page_cache.rst
@@ -0,0 +1,5 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+==========
+Page Cache
+==========
diff --git a/Documentation/mm/page_frags.rst b/Documentation/mm/page_frags.rst
new file mode 100644
index 000000000..7d6f9385d
--- /dev/null
+++ b/Documentation/mm/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/mm/page_migration.rst b/Documentation/mm/page_migration.rst
new file mode 100644
index 000000000..11493bad7
--- /dev/null
+++ b/Documentation/mm/page_migration.rst
@@ -0,0 +1,195 @@
+.. _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.  For compaction purposes, it is also useful to be able to move
+non-LRU pages, such as zsmalloc and virtio-balloon pages.
+
+If a driver wants to make its pages movable, it should define a struct
+movable_operations.  It then needs to call __SetPageMovable() on each
+page that it may be able to move.  This uses the ``page->mapping`` field,
+so this field is not available for the driver to use for other 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 and non-hugetlb page, then
+   this counter is increased by one. If the page was a THP or hugetlb, then
+   this counter is increased by the number of THP or hugetlb 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 or hugetlb.
+
+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.
+
+.. kernel-doc:: include/linux/migrate.h
diff --git a/Documentation/mm/page_owner.rst b/Documentation/mm/page_owner.rst
new file mode 100644
index 000000000..127514955
--- /dev/null
+++ b/Documentation/mm/page_owner.rst
@@ -0,0 +1,189 @@
+.. _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 by default. So, if you'd like to use it, you need
+to add "page_owner=on" to your boot cmdline. If the kernel is built
+with page owner and page owner is disabled in runtime due to not 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.
+
+Although enabling page owner increases kernel size by several kilobytes,
+most of this code is outside page allocator and its hot path. 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 that 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
+
+   The general output of ``page_owner_full.txt`` is as follows::
+
+	Page allocated via order XXX, ...
+	PFN XXX ...
+	// Detailed stack
+
+	Page allocated via order XXX, ...
+	PFN XXX ...
+	// Detailed stack
+    By default, it will do full pfn dump, to start with a given pfn,
+    page_owner supports fseek.
+
+    FILE *fp = fopen("/sys/kernel/debug/page_owner", "r");
+    fseek(fp, pfn_start, SEEK_SET);
+
+   The ``page_owner_sort`` tool ignores ``PFN`` rows, puts the remaining rows
+   in buf, uses regexp to extract the page order value, counts the times
+   and pages of buf, and finally sorts them according to the parameter(s).
+
+   See the result about who allocated each page
+   in the ``sorted_page_owner.txt``. General output::
+
+	XXX times, XXX pages:
+	Page allocated via order XXX, ...
+	// Detailed stack
+
+   By default, ``page_owner_sort`` is sorted according to the times of buf.
+   If you want to sort by the page nums of buf, use the ``-m`` parameter.
+   The detailed parameters are:
+
+   fundamental function::
+
+	Sort:
+		-a		Sort by memory allocation time.
+		-m		Sort by total memory.
+		-p		Sort by pid.
+		-P		Sort by tgid.
+		-n		Sort by task command name.
+		-r		Sort by memory release time.
+		-s		Sort by stack trace.
+		-t		Sort by times (default).
+		--sort <order>	Specify sorting order.  Sorting syntax is [+|-]key[,[+|-]key[,...]].
+				Choose a key from the **STANDARD FORMAT SPECIFIERS** section. The "+" is
+				optional since default direction is increasing numerical or lexicographic
+				order. Mixed use of abbreviated and complete-form of keys is allowed.
+
+		Examples:
+				./page_owner_sort <input> <output> --sort=n,+pid,-tgid
+				./page_owner_sort <input> <output> --sort=at
+
+   additional function::
+
+	Cull:
+		--cull <rules>
+				Specify culling rules.Culling syntax is key[,key[,...]].Choose a
+				multi-letter key from the **STANDARD FORMAT SPECIFIERS** section.
+
+		<rules> is a single argument in the form of a comma-separated list,
+		which offers a way to specify individual culling rules.  The recognized
+		keywords are described in the **STANDARD FORMAT SPECIFIERS** section below.
+		<rules> can be specified by the sequence of keys k1,k2, ..., as described in
+		the STANDARD SORT KEYS section below. Mixed use of abbreviated and
+		complete-form of keys is allowed.
+
+		Examples:
+				./page_owner_sort <input> <output> --cull=stacktrace
+				./page_owner_sort <input> <output> --cull=st,pid,name
+				./page_owner_sort <input> <output> --cull=n,f
+
+	Filter:
+		-f		Filter out the information of blocks whose memory has been released.
+
+	Select:
+		--pid <pidlist>		Select by pid. This selects the blocks whose process ID
+					numbers appear in <pidlist>.
+		--tgid <tgidlist>	Select by tgid. This selects the blocks whose thread
+					group ID numbers appear in <tgidlist>.
+		--name <cmdlist>	Select by task command name. This selects the blocks whose
+					task command name appear in <cmdlist>.
+
+		<pidlist>, <tgidlist>, <cmdlist> are single arguments in the form of a comma-separated list,
+		which offers a way to specify individual selecting rules.
+
+
+		Examples:
+				./page_owner_sort <input> <output> --pid=1
+				./page_owner_sort <input> <output> --tgid=1,2,3
+				./page_owner_sort <input> <output> --name name1,name2
+
+STANDARD FORMAT SPECIFIERS
+==========================
+::
+
+  For --sort option:
+
+	KEY		LONG		DESCRIPTION
+	p		pid		process ID
+	tg		tgid		thread group ID
+	n		name		task command name
+	st		stacktrace	stack trace of the page allocation
+	T		txt		full text of block
+	ft		free_ts		timestamp of the page when it was released
+	at		alloc_ts	timestamp of the page when it was allocated
+	ator		allocator	memory allocator for pages
+
+  For --curl option:
+
+	KEY		LONG		DESCRIPTION
+	p		pid		process ID
+	tg		tgid		thread group ID
+	n		name		task command name
+	f		free		whether the page has been released or not
+	st		stacktrace	stack trace of the page allocation
+	ator		allocator	memory allocator for pages
diff --git a/Documentation/mm/page_reclaim.rst b/Documentation/mm/page_reclaim.rst
new file mode 100644
index 000000000..50a30b7f8
--- /dev/null
+++ b/Documentation/mm/page_reclaim.rst
@@ -0,0 +1,5 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+============
+Page Reclaim
+============
diff --git a/Documentation/mm/page_table_check.rst b/Documentation/mm/page_table_check.rst
new file mode 100644
index 000000000..d19ca356b
--- /dev/null
+++ b/Documentation/mm/page_table_check.rst
@@ -0,0 +1,75 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+.. _page_table_check:
+
+================
+Page Table Check
+================
+
+Introduction
+============
+
+Page table check allows to harden the kernel by ensuring that some types of
+the memory corruptions are prevented.
+
+Page table check performs extra verifications at the time when new pages become
+accessible from the userspace by getting their page table entries (PTEs PMDs
+etc.) added into the table.
+
+In case of detected corruption, the kernel is crashed. There is a small
+performance and memory overhead associated with the page table check. Therefore,
+it is disabled by default, but can be optionally enabled on systems where the
+extra hardening outweighs the performance costs. Also, because page table check
+is synchronous, it can help with debugging double map memory corruption issues,
+by crashing kernel at the time wrong mapping occurs instead of later which is
+often the case with memory corruptions bugs.
+
+Double mapping detection logic
+==============================
+
++-------------------+-------------------+-------------------+------------------+
+| Current Mapping   | New mapping       | Permissions       | Rule             |
++===================+===================+===================+==================+
+| Anonymous         | Anonymous         | Read              | Allow            |
++-------------------+-------------------+-------------------+------------------+
+| Anonymous         | Anonymous         | Read / Write      | Prohibit         |
++-------------------+-------------------+-------------------+------------------+
+| Anonymous         | Named             | Any               | Prohibit         |
++-------------------+-------------------+-------------------+------------------+
+| Named             | Anonymous         | Any               | Prohibit         |
++-------------------+-------------------+-------------------+------------------+
+| Named             | Named             | Any               | Allow            |
++-------------------+-------------------+-------------------+------------------+
+
+Enabling Page Table Check
+=========================
+
+Build kernel with:
+
+- PAGE_TABLE_CHECK=y
+  Note, it can only be enabled on platforms where ARCH_SUPPORTS_PAGE_TABLE_CHECK
+  is available.
+
+- Boot with 'page_table_check=on' kernel parameter.
+
+Optionally, build kernel with PAGE_TABLE_CHECK_ENFORCED in order to have page
+table support without extra kernel parameter.
+
+Implementation notes
+====================
+
+We specifically decided not to use VMA information in order to avoid relying on
+MM states (except for limited "struct page" info). The page table check is a
+separate from Linux-MM state machine that verifies that the user accessible
+pages are not falsely shared.
+
+PAGE_TABLE_CHECK depends on EXCLUSIVE_SYSTEM_RAM. The reason is that without
+EXCLUSIVE_SYSTEM_RAM, users are allowed to map arbitrary physical memory
+regions into the userspace via /dev/mem. At the same time, pages may change
+their properties (e.g., from anonymous pages to named pages) while they are
+still being mapped in the userspace, leading to "corruption" detected by the
+page table check.
+
+Even with EXCLUSIVE_SYSTEM_RAM, I/O pages may be still allowed to be mapped via
+/dev/mem. However, these pages are always considered as named pages, so they
+won't break the logic used in the page table check.
diff --git a/Documentation/mm/page_tables.rst b/Documentation/mm/page_tables.rst
new file mode 100644
index 000000000..96939571d
--- /dev/null
+++ b/Documentation/mm/page_tables.rst
@@ -0,0 +1,5 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+===========
+Page Tables
+===========
diff --git a/Documentation/mm/physical_memory.rst b/Documentation/mm/physical_memory.rst
new file mode 100644
index 000000000..2ab7b8c1c
--- /dev/null
+++ b/Documentation/mm/physical_memory.rst
@@ -0,0 +1,5 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+===============
+Physical Memory
+===============
diff --git a/Documentation/mm/process_addrs.rst b/Documentation/mm/process_addrs.rst
new file mode 100644
index 000000000..e8618fbc6
--- /dev/null
+++ b/Documentation/mm/process_addrs.rst
@@ -0,0 +1,5 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=================
+Process Addresses
+=================
diff --git a/Documentation/mm/remap_file_pages.rst b/Documentation/mm/remap_file_pages.rst
new file mode 100644
index 000000000..7bef6718e
--- /dev/null
+++ b/Documentation/mm/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/mm/shmfs.rst b/Documentation/mm/shmfs.rst
new file mode 100644
index 000000000..8b01ebb4c
--- /dev/null
+++ b/Documentation/mm/shmfs.rst
@@ -0,0 +1,5 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+========================
+Shared Memory Filesystem
+========================
diff --git a/Documentation/mm/slab.rst b/Documentation/mm/slab.rst
new file mode 100644
index 000000000..87d5a5bb1
--- /dev/null
+++ b/Documentation/mm/slab.rst
@@ -0,0 +1,5 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+===============
+Slab Allocation
+===============
diff --git a/Documentation/mm/slub.rst b/Documentation/mm/slub.rst
new file mode 100644
index 000000000..4e1578186
--- /dev/null
+++ b/Documentation/mm/slub.rst
@@ -0,0 +1,461 @@
+.. _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).
+
+
+DebugFS files for SLUB
+======================
+
+For more information about current state of SLUB caches with the user tracking
+debug option enabled, debugfs files are available, typically under
+/sys/kernel/debug/slab/<cache>/ (created only for caches with enabled user
+tracking). There are 2 types of these files with the following debug
+information:
+
+1. alloc_traces::
+
+    Prints information about unique allocation traces of the currently
+    allocated objects. The output is sorted by frequency of each trace.
+
+    Information in the output:
+    Number of objects, allocating function, possible memory wastage of
+    kmalloc objects(total/per-object), minimal/average/maximal jiffies
+    since alloc, pid range of the allocating processes, cpu mask of
+    allocating cpus, numa node mask of origins of memory, and stack trace.
+
+    Example:::
+
+    338 pci_alloc_dev+0x2c/0xa0 waste=521872/1544 age=290837/291891/293509 pid=1 cpus=106 nodes=0-1
+        __kmem_cache_alloc_node+0x11f/0x4e0
+        kmalloc_trace+0x26/0xa0
+        pci_alloc_dev+0x2c/0xa0
+        pci_scan_single_device+0xd2/0x150
+        pci_scan_slot+0xf7/0x2d0
+        pci_scan_child_bus_extend+0x4e/0x360
+        acpi_pci_root_create+0x32e/0x3b0
+        pci_acpi_scan_root+0x2b9/0x2d0
+        acpi_pci_root_add.cold.11+0x110/0xb0a
+        acpi_bus_attach+0x262/0x3f0
+        device_for_each_child+0xb7/0x110
+        acpi_dev_for_each_child+0x77/0xa0
+        acpi_bus_attach+0x108/0x3f0
+        device_for_each_child+0xb7/0x110
+        acpi_dev_for_each_child+0x77/0xa0
+        acpi_bus_attach+0x108/0x3f0
+
+2. free_traces::
+
+    Prints information about unique freeing traces of the currently allocated
+    objects. The freeing traces thus come from the previous life-cycle of the
+    objects and are reported as not available for objects allocated for the first
+    time. The output is sorted by frequency of each trace.
+
+    Information in the output:
+    Number of objects, freeing function, minimal/average/maximal jiffies since free,
+    pid range of the freeing processes, cpu mask of freeing cpus, and stack trace.
+
+    Example:::
+
+    1980 <not-available> age=4294912290 pid=0 cpus=0
+    51 acpi_ut_update_ref_count+0x6a6/0x782 age=236886/237027/237772 pid=1 cpus=1
+	kfree+0x2db/0x420
+	acpi_ut_update_ref_count+0x6a6/0x782
+	acpi_ut_update_object_reference+0x1ad/0x234
+	acpi_ut_remove_reference+0x7d/0x84
+	acpi_rs_get_prt_method_data+0x97/0xd6
+	acpi_get_irq_routing_table+0x82/0xc4
+	acpi_pci_irq_find_prt_entry+0x8e/0x2e0
+	acpi_pci_irq_lookup+0x3a/0x1e0
+	acpi_pci_irq_enable+0x77/0x240
+	pcibios_enable_device+0x39/0x40
+	do_pci_enable_device.part.0+0x5d/0xe0
+	pci_enable_device_flags+0xfc/0x120
+	pci_enable_device+0x13/0x20
+	virtio_pci_probe+0x9e/0x170
+	local_pci_probe+0x48/0x80
+	pci_device_probe+0x105/0x1c0
+
+Christoph Lameter, May 30, 2007
+Sergey Senozhatsky, October 23, 2015
diff --git a/Documentation/mm/split_page_table_lock.rst b/Documentation/mm/split_page_table_lock.rst
new file mode 100644
index 000000000..c08919662
--- /dev/null
+++ b/Documentation/mm/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 are guarded 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/mm/swap.rst b/Documentation/mm/swap.rst
new file mode 100644
index 000000000..78819bd4d
--- /dev/null
+++ b/Documentation/mm/swap.rst
@@ -0,0 +1,5 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+====
+Swap
+====
diff --git a/Documentation/mm/transhuge.rst b/Documentation/mm/transhuge.rst
new file mode 100644
index 000000000..216db1d67
--- /dev/null
+++ b/Documentation/mm/transhuge.rst
@@ -0,0 +1,187 @@
+.. _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.
+
+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/mm/unevictable-lru.rst b/Documentation/mm/unevictable-lru.rst
new file mode 100644
index 000000000..4a0e158aa
--- /dev/null
+++ b/Documentation/mm/unevictable-lru.rst
@@ -0,0 +1,554 @@
+.. _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 node.  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 LRU Page List
+-----------------------------
+
+The Unevictable LRU page list is a lie.  It was never an LRU-ordered list, but a
+companion to the LRU-ordered anonymous and file, active and inactive page lists;
+and now it is not even a page list.  But following familiar convention, here in
+this document and in the source, we often imagine it as a fifth LRU page list.
+
+The Unevictable LRU infrastructure consists of an additional, per-node, 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 as if they were
+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 (or "Movable" pages: outside of consideration here).  If we were to
+     maintain pages elsewhere than on an LRU-like list, where they can be
+     detected by isolate_lru_page(), we would prevent their migration.
+
+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-node LRU
+lists and statistics originally proposed and posted by Christoph Lameter.
+
+
+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-node unevictable
+list as a result of the "arrayification" of the per-node 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 mm/internal.h 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 (e.g. 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 (e.g. 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 an 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 memory cgroup and node 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_active_list() or shrink_page_list() where they will be detected when
+vmscan walks the reverse map in folio_referenced() or try_to_unmap().  The page
+is culled to the unevictable list when it is released by the shrinker.
+
+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, __pagevec_lru_add_fn() will recheck the
+unevictable state of a page before placing it on the unevictable list.
+
+
+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 (Rik van Riel had the same idea three years
+earlier).  But 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 detect 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 when munlocking, to determine whether any
+other VM_LOCKED VMAs still mapped the page.
+
+However, walking the reverse map for each page when munlocking was ugly and
+inefficient, and could lead to catastrophic contention on a file's rmap lock,
+when many processes which had it mlocked were trying to exit.  In 5.18, the
+idea of keeping mlock_count in Unevictable LRU list link field was revived and
+put to work, without preventing the migration of mlocked pages.  This is why
+the "Unevictable LRU list" cannot be a linked list of pages now; but there was
+no use for that linked list anyway - though its size is maintained for meminfo.
+
+
+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()/mlock2()/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 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 by folio_referenced() or try_to_unmap().
+
+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()/mlock2()/mlockall() System Call Handling
+------------------------------------------------
+
+mlock(), mlock2() and mlockall() system call handlers proceed to 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.  Any pages
+already present in the VMA are then marked as mlocked by mlock_page() via
+mlock_pte_range() via walk_page_range() via mlock_vma_pages_range().
+
+Before returning from the system call, do_mlock() or mlockall() will call
+__mm_populate() to fault in the remaining pages via get_user_pages() and to
+mark those pages as mlocked as they are faulted.
+
+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, they will be handled in the
+fault path - which is also how mlock2()'s MLOCK_ONFAULT areas are handled.
+
+For each PTE (or PMD) being faulted into a VMA, the page add rmap function
+calls mlock_vma_page(), which calls mlock_page() when the VMA is VM_LOCKED
+(unless it is a PTE mapping of a part of a transparent huge page).  Or when
+it is a newly allocated anonymous page, lru_cache_add_inactive_or_unevictable()
+calls mlock_new_page() instead: similar to mlock_page(), but can make better
+judgments, since this page is held exclusively and known not to be on LRU yet.
+
+mlock_page() sets PageMlocked immediately, then places the page on the CPU's
+mlock pagevec, to batch up the rest of the work to be done under lru_lock by
+__mlock_page().  __mlock_page() sets PageUnevictable, initializes mlock_count
+and moves the page to unevictable state ("the unevictable LRU", but with
+mlock_count in place of LRU threading).  Or if the page was already PageLRU
+and PageUnevictable and PageMlocked, it simply increments the mlock_count.
+
+But in practice that may not work ideally: the page may not yet be on an LRU, or
+it may have been temporarily isolated from LRU.  In such cases the mlock_count
+field cannot be touched, but will be set to 0 later when __pagevec_lru_add_fn()
+returns the page to "LRU".  Races prohibit mlock_count from being set to 1 then:
+rather than risk stranding a page indefinitely as unevictable, always err with
+mlock_count on the low side, so that when munlocked the page will be rescued to
+an evictable LRU, then perhaps be mlocked again later if vmscan finds it in a
+VM_LOCKED VMA.
+
+
+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.  But __mm_populate() includes
+   hugetlbfs ranges, allocating the huge pages and populating 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.  __mm_populate() includes
+   these ranges, populating the PTEs if not already populated.
+
+4) VMAs with VM_MIXEDMAP set are not marked VM_LOCKED, but __mm_populate()
+   includes these ranges, populating the PTEs if not already populated.
+
+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() System Call Handling
+-------------------------------------------
+
+The munlock() and munlockall() system calls are handled by the same
+mlock_fixup() function as mlock(), mlock2() and mlockall() system calls are.
+If called to munlock 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, those VMAs will be ignored for munlock.
+
+If the VMA is VM_LOCKED, mlock_fixup() again attempts to merge or split off the
+specified range.  All pages in the VMA are then munlocked by munlock_page() via
+mlock_pte_range() via walk_page_range() via mlock_vma_pages_range() - the same
+function used when mlocking a VMA range, with new flags for the VMA indicating
+that it is munlock() being performed.
+
+munlock_page() uses the mlock pagevec to batch up work to be done under
+lru_lock by  __munlock_page().  __munlock_page() decrements the page's
+mlock_count, and when that reaches 0 it clears PageMlocked and clears
+PageUnevictable, moving the page from unevictable state to inactive LRU.
+
+But in practice that may not work ideally: the page may not yet have reached
+"the unevictable LRU", or it may have been temporarily isolated from it.  In
+those cases its mlock_count field is unusable and must be assumed to be 0: so
+that the page will be rescued to an evictable LRU, then perhaps be mlocked
+again later if vmscan finds it in a VM_LOCKED VMA.
+
+
+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.  PG_mlocked is cleared from the
+the old page when it is unmapped from the last VM_LOCKED VMA, and set when the
+new page is mapped in place of migration entry in a VM_LOCKED VMA.  If the page
+was unevictable because mlocked, PG_unevictable follows PG_mlocked; but if the
+page was unevictable for other reasons, PG_unevictable is copied explicitly.
+
+Note that page migration can race with mlocking or munlocking of the same page.
+There is mostly no problem since page migration requires unmapping all PTEs of
+the old page (including munlock where VM_LOCKED), then mapping in the new page
+(including mlock where VM_LOCKED).  The page table locks provide sufficient
+synchronization.
+
+However, since mlock_vma_pages_range() starts by setting VM_LOCKED on a VMA,
+before mlocking any pages already present, if one of those pages were migrated
+before mlock_pte_range() reached it, it would get counted twice in mlock_count.
+To prevent that, mlock_vma_pages_range() temporarily marks the VMA as VM_IO,
+so that mlock_vma_page() will skip it.
+
+To complete page migration, we place the old and new pages back onto the LRU
+afterwards.  The "unneeded" page - old page on success, new page on failure -
+is freed when the reference count held by the migration process is released.
+
+
+Compacting MLOCKED Pages
+------------------------
+
+The memory map can be scanned for compactable regions and the default behavior
+is to let unevictable pages be moved.  /proc/sys/vm/compact_unevictable_allowed
+controls this behavior (see Documentation/admin-guide/sysctl/vm.rst).  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, and no user mlock()s the
+whole of the 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 a new intermittent failure mode for the syscall is undesirable.
+
+We handle this by keeping PTE-mlocked huge pages on evictable LRU lists:
+the PMD on the border of a VM_LOCKED VMA will be split into a 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 the unevictable LRU and the rest can be reclaimed.
+
+/proc/meminfo's Unevictable and Mlocked amounts do not include those parts
+of a transparent huge page which are mapped only by PTEs in VM_LOCKED VMAs.
+
+
+mmap(MAP_LOCKED) System Call Handling
+-------------------------------------
+
+In addition to the mlock(), mlock2() and mlockall() system calls, an application
+can request that a region of memory be mlocked by 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 - 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.
+
+
+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.
+
+For each PTE (or PMD) being unmapped from a VMA, page_remove_rmap() calls
+munlock_vma_page(), which calls munlock_page() when the VMA is VM_LOCKED
+(unless it was a PTE mapping of a part of a transparent huge page).
+
+munlock_page() uses the mlock pagevec to batch up work to be done under
+lru_lock by  __munlock_page().  __munlock_page() decrements the page's
+mlock_count, and when that reaches 0 it clears PageMlocked and clears
+PageUnevictable, moving the page from unevictable state to inactive LRU.
+
+But in practice that may not work ideally: the page may not yet have reached
+"the unevictable LRU", or it may have been temporarily isolated from it.  In
+those cases its mlock_count field is unusable and must be assumed to be 0: so
+that the page will be rescued to an evictable LRU, then perhaps be mlocked
+again later if vmscan finds it in a VM_LOCKED VMA.
+
+
+Truncating MLOCKED Pages
+------------------------
+
+File truncation or hole punching forcibly unmaps the deleted pages from
+userspace; truncation even unmaps and deletes any private anonymous pages
+which had been Copied-On-Write from the file pages now being truncated.
+
+Mlocked pages can be munlocked and deleted in this way: like with munmap(),
+for each PTE (or PMD) being unmapped from a VMA, page_remove_rmap() calls
+munlock_vma_page(), which calls munlock_page() when the VMA is VM_LOCKED
+(unless it was a PTE mapping of a part of a transparent huge page).
+
+However, if there is a racing munlock(), since mlock_vma_pages_range() starts
+munlocking by clearing VM_LOCKED from a VMA, before munlocking all the pages
+present, if one of those pages were unmapped by truncation or hole punch before
+mlock_pte_range() reached it, it would not be recognized as mlocked by this VMA,
+and would not be counted out of mlock_count.  In this rare case, a page may
+still appear as PageMlocked after it has been fully unmapped: and it is left to
+release_pages() (or __page_cache_release()) to clear it and update statistics
+before freeing (this event is counted in /proc/vmstat unevictable_pgs_cleared,
+which is usually 0).
+
+
+Page Reclaim in shrink_*_list()
+-------------------------------
+
+vmscan's shrink_active_list() culls any obviously unevictable pages -
+i.e. !page_evictable(page) pages - diverting those 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) pages still mapped into VM_LOCKED VMAs, which should be marked mlocked,
+     but events left mlock_count too low, so they were munlocked too early.
+
+vmscan's shrink_inactive_list() and shrink_page_list() also divert obviously
+unevictable pages found on the inactive lists to the appropriate memory cgroup
+and node unevictable list.
+
+rmap's folio_referenced_one(), called via vmscan's shrink_active_list() or
+shrink_page_list(), and rmap's try_to_unmap_one() called via shrink_page_list(),
+check for (3) pages still mapped into VM_LOCKED VMAs, and call mlock_vma_page()
+to correct them.  Such pages are culled to the unevictable list when released
+by the shrinker.
diff --git a/Documentation/mm/vmalloc.rst b/Documentation/mm/vmalloc.rst
new file mode 100644
index 000000000..363fe20d6
--- /dev/null
+++ b/Documentation/mm/vmalloc.rst
@@ -0,0 +1,5 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+======================================
+Virtually Contiguous Memory Allocation
+======================================
diff --git a/Documentation/mm/vmalloced-kernel-stacks.rst b/Documentation/mm/vmalloced-kernel-stacks.rst
new file mode 100644
index 000000000..fc8c67833
--- /dev/null
+++ b/Documentation/mm/vmalloced-kernel-stacks.rst
@@ -0,0 +1,153 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=====================================
+Virtually Mapped Kernel Stack Support
+=====================================
+
+:Author: Shuah Khan <skhan@linuxfoundation.org>
+
+.. contents:: :local:
+
+Overview
+--------
+
+This is a compilation of information from the code and original patch
+series that introduced the `Virtually Mapped Kernel Stacks feature
+<https://lwn.net/Articles/694348/>`
+
+Introduction
+------------
+
+Kernel stack overflows are often hard to debug and make the kernel
+susceptible to exploits. Problems could show up at a later time making
+it difficult to isolate and root-cause.
+
+Virtually-mapped kernel stacks with guard pages causes kernel stack
+overflows to be caught immediately rather than causing difficult to
+diagnose corruptions.
+
+HAVE_ARCH_VMAP_STACK and VMAP_STACK configuration options enable
+support for virtually mapped stacks with guard pages. This feature
+causes reliable faults when the stack overflows. The usability of
+the stack trace after overflow and response to the overflow itself
+is architecture dependent.
+
+.. note::
+        As of this writing, arm64, powerpc, riscv, s390, um, and x86 have
+        support for VMAP_STACK.
+
+HAVE_ARCH_VMAP_STACK
+--------------------
+
+Architectures that can support Virtually Mapped Kernel Stacks should
+enable this bool configuration option. The requirements are:
+
+- vmalloc space must be large enough to hold many kernel stacks. This
+  may rule out many 32-bit architectures.
+- Stacks in vmalloc space need to work reliably.  For example, if
+  vmap page tables are created on demand, either this mechanism
+  needs to work while the stack points to a virtual address with
+  unpopulated page tables or arch code (switch_to() and switch_mm(),
+  most likely) needs to ensure that the stack's page table entries
+  are populated before running on a possibly unpopulated stack.
+- If the stack overflows into a guard page, something reasonable
+  should happen. The definition of "reasonable" is flexible, but
+  instantly rebooting without logging anything would be unfriendly.
+
+VMAP_STACK
+----------
+
+VMAP_STACK bool configuration option when enabled allocates virtually
+mapped task stacks. This option depends on HAVE_ARCH_VMAP_STACK.
+
+- Enable this if you want the use virtually-mapped kernel stacks
+  with guard pages. This causes kernel stack overflows to be caught
+  immediately rather than causing difficult-to-diagnose corruption.
+
+.. note::
+
+        Using this feature with KASAN requires architecture support
+        for backing virtual mappings with real shadow memory, and
+        KASAN_VMALLOC must be enabled.
+
+.. note::
+
+        VMAP_STACK is enabled, it is not possible to run DMA on stack
+        allocated data.
+
+Kernel configuration options and dependencies keep changing. Refer to
+the latest code base:
+
+`Kconfig <https://git.kernel.org/pub/scm/linux/kernel/git/torvalds/linux.git/tree/arch/Kconfig>`
+
+Allocation
+-----------
+
+When a new kernel thread is created, thread stack is allocated from
+virtually contiguous memory pages from the page level allocator. These
+pages are mapped into contiguous kernel virtual space with PAGE_KERNEL
+protections.
+
+alloc_thread_stack_node() calls __vmalloc_node_range() to allocate stack
+with PAGE_KERNEL protections.
+
+- Allocated stacks are cached and later reused by new threads, so memcg
+  accounting is performed manually on assigning/releasing stacks to tasks.
+  Hence, __vmalloc_node_range is called without __GFP_ACCOUNT.
+- vm_struct is cached to be able to find when thread free is initiated
+  in interrupt context. free_thread_stack() can be called in interrupt
+  context.
+- On arm64, all VMAP's stacks need to have the same alignment to ensure
+  that VMAP'd stack overflow detection works correctly. Arch specific
+  vmap stack allocator takes care of this detail.
+- This does not address interrupt stacks - according to the original patch
+
+Thread stack allocation is initiated from clone(), fork(), vfork(),
+kernel_thread() via kernel_clone(). Leaving a few hints for searching
+the code base to understand when and how thread stack is allocated.
+
+Bulk of the code is in:
+`kernel/fork.c <https://git.kernel.org/pub/scm/linux/kernel/git/torvalds/linux.git/tree/kernel/fork.c>`.
+
+stack_vm_area pointer in task_struct keeps track of the virtually allocated
+stack and a non-null stack_vm_area pointer serves as a indication that the
+virtually mapped kernel stacks are enabled.
+
+::
+
+        struct vm_struct *stack_vm_area;
+
+Stack overflow handling
+-----------------------
+
+Leading and trailing guard pages help detect stack overflows. When stack
+overflows into the guard pages, handlers have to be careful not overflow
+the stack again. When handlers are called, it is likely that very little
+stack space is left.
+
+On x86, this is done by handling the page fault indicating the kernel
+stack overflow on the double-fault stack.
+
+Testing VMAP allocation with guard pages
+----------------------------------------
+
+How do we ensure that VMAP_STACK is actually allocating with a leading
+and trailing guard page? The following lkdtm tests can help detect any
+regressions.
+
+::
+
+        void lkdtm_STACK_GUARD_PAGE_LEADING()
+        void lkdtm_STACK_GUARD_PAGE_TRAILING()
+
+Conclusions
+-----------
+
+- A percpu cache of vmalloced stacks appears to be a bit faster than a
+  high-order stack allocation, at least when the cache hits.
+- THREAD_INFO_IN_TASK gets rid of arch-specific thread_info entirely and
+  simply embed the thread_info (containing only flags) and 'int cpu' into
+  task_struct.
+- The thread stack can be free'ed as soon as the task is dead (without
+  waiting for RCU) and then, if vmapped stacks are in use, cache the
+  entire stack for reuse on the same cpu.
diff --git a/Documentation/mm/vmemmap_dedup.rst b/Documentation/mm/vmemmap_dedup.rst
new file mode 100644
index 000000000..a4b12ff90
--- /dev/null
+++ b/Documentation/mm/vmemmap_dedup.rst
@@ -0,0 +1,249 @@
+.. SPDX-License-Identifier: GPL-2.0
+
+=========================================
+A vmemmap diet for HugeTLB and Device DAX
+=========================================
+
+HugeTLB
+=======
+
+This section is to explain how HugeTLB Vmemmap Optimization (HVO) works.
+
+The ``struct page`` structures are used to describe a physical page frame. By
+default, there is a one-to-one mapping from a page frame to it's corresponding
+``struct page``.
+
+HugeTLB pages consist of multiple base page size pages and is supported by many
+architectures. See Documentation/admin-guide/mm/hugetlbpage.rst for more
+details. On the x86-64 architecture, HugeTLB pages of size 2MB and 1GB are
+currently supported. Since the base page size on x86 is 4KB, a 2MB HugeTLB page
+consists of 512 base pages and a 1GB HugeTLB page consists of 4096 base pages.
+For each base page, there is a corresponding ``struct page``.
+
+Within the HugeTLB subsystem, only the first 4 ``struct page`` are used to
+contain unique information about a HugeTLB page. ``__NR_USED_SUBPAGE`` provides
+this upper limit. The only 'useful' information in the remaining ``struct page``
+is the compound_head field, and this field is the same for all tail pages.
+
+By removing redundant ``struct page`` for HugeTLB pages, memory can be returned
+to the buddy allocator for other uses.
+
+Different architectures support different HugeTLB pages. For example, the
+following table is the HugeTLB page size supported by x86 and arm64
+architectures. Because arm64 supports 4k, 16k, and 64k base pages and
+supports contiguous entries, so it supports many kinds of sizes of HugeTLB
+page.
+
++--------------+-----------+-----------------------------------------------+
+| Architecture | Page Size |                HugeTLB Page Size              |
++--------------+-----------+-----------+-----------+-----------+-----------+
+|    x86-64    |    4KB    |    2MB    |    1GB    |           |           |
++--------------+-----------+-----------+-----------+-----------+-----------+
+|              |    4KB    |   64KB    |    2MB    |    32MB   |    1GB    |
+|              +-----------+-----------+-----------+-----------+-----------+
+|    arm64     |   16KB    |    2MB    |   32MB    |     1GB   |           |
+|              +-----------+-----------+-----------+-----------+-----------+
+|              |   64KB    |    2MB    |  512MB    |    16GB   |           |
++--------------+-----------+-----------+-----------+-----------+-----------+
+
+When the system boot up, every HugeTLB page has more than one ``struct page``
+structs which size is (unit: pages)::
+
+   struct_size = HugeTLB_Size / PAGE_SIZE * sizeof(struct page) / PAGE_SIZE
+
+Where HugeTLB_Size is the size of the HugeTLB page. We know that the size
+of the HugeTLB page is always n times PAGE_SIZE. So we can get the following
+relationship::
+
+   HugeTLB_Size = n * PAGE_SIZE
+
+Then::
+
+   struct_size = n * PAGE_SIZE / PAGE_SIZE * sizeof(struct page) / PAGE_SIZE
+               = n * sizeof(struct page) / PAGE_SIZE
+
+We can use huge mapping at the pud/pmd level for the HugeTLB page.
+
+For the HugeTLB page of the pmd level mapping, then::
+
+   struct_size = n * sizeof(struct page) / PAGE_SIZE
+               = PAGE_SIZE / sizeof(pte_t) * sizeof(struct page) / PAGE_SIZE
+               = sizeof(struct page) / sizeof(pte_t)
+               = 64 / 8
+               = 8 (pages)
+
+Where n is how many pte entries which one page can contains. So the value of
+n is (PAGE_SIZE / sizeof(pte_t)).
+
+This optimization only supports 64-bit system, so the value of sizeof(pte_t)
+is 8. And this optimization also applicable only when the size of ``struct page``
+is a power of two. In most cases, the size of ``struct page`` is 64 bytes (e.g.
+x86-64 and arm64). So if we use pmd level mapping for a HugeTLB page, the
+size of ``struct page`` structs of it is 8 page frames which size depends on the
+size of the base page.
+
+For the HugeTLB page of the pud level mapping, then::
+
+   struct_size = PAGE_SIZE / sizeof(pmd_t) * struct_size(pmd)
+               = PAGE_SIZE / 8 * 8 (pages)
+               = PAGE_SIZE (pages)
+
+Where the struct_size(pmd) is the size of the ``struct page`` structs of a
+HugeTLB page of the pmd level mapping.
+
+E.g.: A 2MB HugeTLB page on x86_64 consists in 8 page frames while 1GB
+HugeTLB page consists in 4096.
+
+Next, we take the pmd level mapping of the HugeTLB page as an example to
+show the internal implementation of this optimization. There are 8 pages
+``struct page`` structs associated with a HugeTLB page which is pmd mapped.
+
+Here is how things look before optimization::
+
+    HugeTLB                  struct pages(8 pages)         page frame(8 pages)
+ +-----------+ ---virt_to_page---> +-----------+   mapping to   +-----------+
+ |           |                     |     0     | -------------> |     0     |
+ |           |                     +-----------+                +-----------+
+ |           |                     |     1     | -------------> |     1     |
+ |           |                     +-----------+                +-----------+
+ |           |                     |     2     | -------------> |     2     |
+ |           |                     +-----------+                +-----------+
+ |           |                     |     3     | -------------> |     3     |
+ |           |                     +-----------+                +-----------+
+ |           |                     |     4     | -------------> |     4     |
+ |    PMD    |                     +-----------+                +-----------+
+ |   level   |                     |     5     | -------------> |     5     |
+ |  mapping  |                     +-----------+                +-----------+
+ |           |                     |     6     | -------------> |     6     |
+ |           |                     +-----------+                +-----------+
+ |           |                     |     7     | -------------> |     7     |
+ |           |                     +-----------+                +-----------+
+ |           |
+ |           |
+ |           |
+ +-----------+
+
+The value of page->compound_head is the same for all tail pages. The first
+page of ``struct page`` (page 0) associated with the HugeTLB page contains the 4
+``struct page`` necessary to describe the HugeTLB. The only use of the remaining
+pages of ``struct page`` (page 1 to page 7) is to point to page->compound_head.
+Therefore, we can remap pages 1 to 7 to page 0. Only 1 page of ``struct page``
+will be used for each HugeTLB page. This will allow us to free the remaining
+7 pages to the buddy allocator.
+
+Here is how things look after remapping::
+
+    HugeTLB                  struct pages(8 pages)         page frame(8 pages)
+ +-----------+ ---virt_to_page---> +-----------+   mapping to   +-----------+
+ |           |                     |     0     | -------------> |     0     |
+ |           |                     +-----------+                +-----------+
+ |           |                     |     1     | ---------------^ ^ ^ ^ ^ ^ ^
+ |           |                     +-----------+                  | | | | | |
+ |           |                     |     2     | -----------------+ | | | | |
+ |           |                     +-----------+                    | | | | |
+ |           |                     |     3     | -------------------+ | | | |
+ |           |                     +-----------+                      | | | |
+ |           |                     |     4     | ---------------------+ | | |
+ |    PMD    |                     +-----------+                        | | |
+ |   level   |                     |     5     | -----------------------+ | |
+ |  mapping  |                     +-----------+                          | |
+ |           |                     |     6     | -------------------------+ |
+ |           |                     +-----------+                            |
+ |           |                     |     7     | ---------------------------+
+ |           |                     +-----------+
+ |           |
+ |           |
+ |           |
+ +-----------+
+
+When a HugeTLB is freed to the buddy system, we should allocate 7 pages for
+vmemmap pages and restore the previous mapping relationship.
+
+For the HugeTLB page of the pud level mapping. It is similar to the former.
+We also can use this approach to free (PAGE_SIZE - 1) vmemmap pages.
+
+Apart from the HugeTLB page of the pmd/pud level mapping, some architectures
+(e.g. aarch64) provides a contiguous bit in the translation table entries
+that hints to the MMU to indicate that it is one of a contiguous set of
+entries that can be cached in a single TLB entry.
+
+The contiguous bit is used to increase the mapping size at the pmd and pte
+(last) level. So this type of HugeTLB page can be optimized only when its
+size of the ``struct page`` structs is greater than **1** page.
+
+Notice: The head vmemmap page is not freed to the buddy allocator and all
+tail vmemmap pages are mapped to the head vmemmap page frame. So we can see
+more than one ``struct page`` struct with ``PG_head`` (e.g. 8 per 2 MB HugeTLB
+page) associated with each HugeTLB page. The ``compound_head()`` can handle
+this correctly. There is only **one** head ``struct page``, the tail
+``struct page`` with ``PG_head`` are fake head ``struct page``.  We need an
+approach to distinguish between those two different types of ``struct page`` so
+that ``compound_head()`` can return the real head ``struct page`` when the
+parameter is the tail ``struct page`` but with ``PG_head``. The following code
+snippet describes how to distinguish between real and fake head ``struct page``.
+
+.. code-block:: c
+
+	if (test_bit(PG_head, &page->flags)) {
+		unsigned long head = READ_ONCE(page[1].compound_head);
+
+		if (head & 1) {
+			if (head == (unsigned long)page + 1)
+				/* head struct page */
+			else
+				/* tail struct page */
+		} else {
+			/* head struct page */
+		}
+	}
+
+We can safely access the field of the **page[1]** with ``PG_head`` because the
+page is a compound page composed with at least two contiguous pages.
+The implementation refers to ``page_fixed_fake_head()``.
+
+Device DAX
+==========
+
+The device-dax interface uses the same tail deduplication technique explained
+in the previous chapter, except when used with the vmemmap in
+the device (altmap).
+
+The following page sizes are supported in DAX: PAGE_SIZE (4K on x86_64),
+PMD_SIZE (2M on x86_64) and PUD_SIZE (1G on x86_64).
+
+The differences with HugeTLB are relatively minor.
+
+It only use 3 ``struct page`` for storing all information as opposed
+to 4 on HugeTLB pages.
+
+There's no remapping of vmemmap given that device-dax memory is not part of
+System RAM ranges initialized at boot. Thus the tail page deduplication
+happens at a later stage when we populate the sections. HugeTLB reuses the
+the head vmemmap page representing, whereas device-dax reuses the tail
+vmemmap page. This results in only half of the savings compared to HugeTLB.
+
+Deduplicated tail pages are not mapped read-only.
+
+Here's how things look like on device-dax after the sections are populated::
+
+ +-----------+ ---virt_to_page---> +-----------+   mapping to   +-----------+
+ |           |                     |     0     | -------------> |     0     |
+ |           |                     +-----------+                +-----------+
+ |           |                     |     1     | -------------> |     1     |
+ |           |                     +-----------+                +-----------+
+ |           |                     |     2     | ----------------^ ^ ^ ^ ^ ^
+ |           |                     +-----------+                   | | | | |
+ |           |                     |     3     | ------------------+ | | | |
+ |           |                     +-----------+                     | | | |
+ |           |                     |     4     | --------------------+ | | |
+ |    PMD    |                     +-----------+                       | | |
+ |   level   |                     |     5     | ----------------------+ | |
+ |  mapping  |                     +-----------+                         | |
+ |           |                     |     6     | ------------------------+ |
+ |           |                     +-----------+                           |
+ |           |                     |     7     | --------------------------+
+ |           |                     +-----------+
+ |           |
+ |           |
+ |           |
+ +-----------+
diff --git a/Documentation/mm/z3fold.rst b/Documentation/mm/z3fold.rst
new file mode 100644
index 000000000..224e3c61d
--- /dev/null
+++ b/Documentation/mm/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/mm/zsmalloc.rst b/Documentation/mm/zsmalloc.rst
new file mode 100644
index 000000000..d98607aa9
--- /dev/null
+++ b/Documentation/mm/zsmalloc.rst
@@ -0,0 +1,84 @@
+.. _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
+freeable
+	the approximate number of pages class compaction can free
+
+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