diff options
Diffstat (limited to 'Documentation/vm')
| -rw-r--r-- | Documentation/vm/00-INDEX | 58 | ||||
| -rw-r--r-- | Documentation/vm/active_mm.rst | 91 | ||||
| -rw-r--r-- | Documentation/vm/active_mm.txt | 83 | ||||
| -rw-r--r-- | Documentation/vm/balance.rst (renamed from Documentation/vm/balance) | 15 | ||||
| -rw-r--r-- | Documentation/vm/cleancache.rst (renamed from Documentation/vm/cleancache.txt) | 105 | ||||
| -rw-r--r-- | Documentation/vm/conf.py | 10 | ||||
| -rw-r--r-- | Documentation/vm/frontswap.rst (renamed from Documentation/vm/frontswap.txt) | 59 | ||||
| -rw-r--r-- | Documentation/vm/highmem.rst (renamed from Documentation/vm/highmem.txt) | 87 | ||||
| -rw-r--r-- | Documentation/vm/hmm.rst (renamed from Documentation/vm/hmm.txt) | 78 | ||||
| -rw-r--r-- | Documentation/vm/hugetlbfs_reserv.rst (renamed from Documentation/vm/hugetlbfs_reserv.txt) | 220 | ||||
| -rw-r--r-- | Documentation/vm/hugetlbpage.txt | 351 | ||||
| -rw-r--r-- | Documentation/vm/hwpoison.rst (renamed from Documentation/vm/hwpoison.txt) | 141 | ||||
| -rw-r--r-- | Documentation/vm/idle_page_tracking.txt | 98 | ||||
| -rw-r--r-- | Documentation/vm/index.rst | 50 | ||||
| -rw-r--r-- | Documentation/vm/ksm.rst | 87 | ||||
| -rw-r--r-- | Documentation/vm/ksm.txt | 178 | ||||
| -rw-r--r-- | Documentation/vm/mmu_notifier.rst | 99 | ||||
| -rw-r--r-- | Documentation/vm/mmu_notifier.txt | 93 | ||||
| -rw-r--r-- | Documentation/vm/numa.rst (renamed from Documentation/vm/numa) | 6 | ||||
| -rw-r--r-- | Documentation/vm/numa_memory_policy.txt | 452 | ||||
| -rw-r--r-- | Documentation/vm/overcommit-accounting | 80 | ||||
| -rw-r--r-- | Documentation/vm/overcommit-accounting.rst | 87 | ||||
| -rw-r--r-- | Documentation/vm/page_frags.rst (renamed from Documentation/vm/page_frags) | 5 | ||||
| -rw-r--r-- | Documentation/vm/page_migration.rst (renamed from Documentation/vm/page_migration) | 149 | ||||
| -rw-r--r-- | Documentation/vm/page_owner.rst (renamed from Documentation/vm/page_owner.txt) | 34 | ||||
| -rw-r--r-- | Documentation/vm/pagemap.txt | 183 | ||||
| -rw-r--r-- | Documentation/vm/remap_file_pages.rst (renamed from Documentation/vm/remap_file_pages.txt) | 6 | ||||
| -rw-r--r-- | Documentation/vm/slub.rst | 361 | ||||
| -rw-r--r-- | Documentation/vm/slub.txt | 342 | ||||
| -rw-r--r-- | Documentation/vm/soft-dirty.txt | 43 | ||||
| -rw-r--r-- | Documentation/vm/split_page_table_lock.rst (renamed from Documentation/vm/split_page_table_lock) | 12 | ||||
| -rw-r--r-- | Documentation/vm/swap_numa.rst (renamed from Documentation/vm/swap_numa.txt) | 55 | ||||
| -rw-r--r-- | Documentation/vm/transhuge.rst | 197 | ||||
| -rw-r--r-- | Documentation/vm/transhuge.txt | 527 | ||||
| -rw-r--r-- | Documentation/vm/unevictable-lru.rst (renamed from Documentation/vm/unevictable-lru.txt) | 117 | ||||
| -rw-r--r-- | Documentation/vm/userfaultfd.txt | 229 | ||||
| -rw-r--r-- | Documentation/vm/z3fold.rst (renamed from Documentation/vm/z3fold.txt) | 6 | ||||
| -rw-r--r-- | Documentation/vm/zsmalloc.rst (renamed from Documentation/vm/zsmalloc.txt) | 60 | ||||
| -rw-r--r-- | Documentation/vm/zswap.rst (renamed from Documentation/vm/zswap.txt) | 71 |
39 files changed, 1685 insertions, 3240 deletions
diff --git a/Documentation/vm/00-INDEX b/Documentation/vm/00-INDEX index 0278f2c85efb5..f4a4f3e884cf6 100644 --- a/Documentation/vm/00-INDEX +++ b/Documentation/vm/00-INDEX @@ -1,62 +1,50 @@ 00-INDEX - this file. -active_mm.txt +active_mm.rst - An explanation from Linus about tsk->active_mm vs tsk->mm. -balance +balance.rst - various information on memory balancing. -cleancache.txt +cleancache.rst - Intro to cleancache and page-granularity victim cache. -frontswap.txt +frontswap.rst - Outline frontswap, part of the transcendent memory frontend. -highmem.txt +highmem.rst - Outline of highmem and common issues. -hmm.txt +hmm.rst - Documentation of heterogeneous memory management -hugetlbpage.txt - - a brief summary of hugetlbpage support in the Linux kernel. -hugetlbfs_reserv.txt +hugetlbfs_reserv.rst - A brief overview of hugetlbfs reservation design/implementation. -hwpoison.txt +hwpoison.rst - explains what hwpoison is -idle_page_tracking.txt - - description of the idle page tracking feature. -ksm.txt +ksm.rst - how to use the Kernel Samepage Merging feature. -mmu_notifier.txt +mmu_notifier.rst - a note about clearing pte/pmd and mmu notifications -numa +numa.rst - information about NUMA specific code in the Linux vm. -numa_memory_policy.txt - - documentation of concepts and APIs of the 2.6 memory policy support. -overcommit-accounting +overcommit-accounting.rst - description of the Linux kernels overcommit handling modes. -page_frags +page_frags.rst - description of page fragments allocator -page_migration +page_migration.rst - description of page migration in NUMA systems. -pagemap.txt - - pagemap, from the userspace perspective -page_owner.txt +page_owner.rst - tracking about who allocated each page -remap_file_pages.txt +remap_file_pages.rst - a note about remap_file_pages() system call -slub.txt +slub.rst - a short users guide for SLUB. -soft-dirty.txt - - short explanation for soft-dirty PTEs -split_page_table_lock +split_page_table_lock.rst - Separate per-table lock to improve scalability of the old page_table_lock. -swap_numa.txt +swap_numa.rst - automatic binding of swap device to numa node -transhuge.txt +transhuge.rst - Transparent Hugepage Support, alternative way of using hugepages. -unevictable-lru.txt +unevictable-lru.rst - Unevictable LRU infrastructure -userfaultfd.txt - - description of userfaultfd system call z3fold.txt - outline of z3fold allocator for storing compressed pages -zsmalloc.txt +zsmalloc.rst - outline of zsmalloc allocator for storing compressed pages -zswap.txt +zswap.rst - Intro to compressed cache for swap pages diff --git a/Documentation/vm/active_mm.rst b/Documentation/vm/active_mm.rst new file mode 100644 index 0000000000000..c84471b180f8b --- /dev/null +++ b/Documentation/vm/active_mm.rst @@ -0,0 +1,91 @@ +.. _active_mm: + +========= +Active MM +========= + +:: + + List: linux-kernel + Subject: Re: active_mm + From: Linus Torvalds <torvalds () transmeta ! com> + Date: 1999-07-30 21:36:24 + + Cc'd to linux-kernel, because I don't write explanations all that often, + and when I do I feel better about more people reading them. + + On Fri, 30 Jul 1999, David Mosberger wrote: + > + > Is there a brief description someplace on how "mm" vs. "active_mm" in + > the task_struct are supposed to be used? (My apologies if this was + > discussed on the mailing lists---I just returned from vacation and + > wasn't able to follow linux-kernel for a while). + + Basically, the new setup is: + + - we have "real address spaces" and "anonymous address spaces". The + difference is that an anonymous address space doesn't care about the + user-level page tables at all, so when we do a context switch into an + anonymous address space we just leave the previous address space + active. + + The obvious use for a "anonymous address space" is any thread that + doesn't need any user mappings - all kernel threads basically fall into + this category, but even "real" threads can temporarily say that for + some amount of time they are not going to be interested in user space, + and that the scheduler might as well try to avoid wasting time on + switching the VM state around. Currently only the old-style bdflush + sync does that. + + - "tsk->mm" points to the "real address space". For an anonymous process, + tsk->mm will be NULL, for the logical reason that an anonymous process + really doesn't _have_ a real address space at all. + + - however, we obviously need to keep track of which address space we + "stole" for such an anonymous user. For that, we have "tsk->active_mm", + which shows what the currently active address space is. + + The rule is that for a process with a real address space (ie tsk->mm is + non-NULL) the active_mm obviously always has to be the same as the real + one. + + For a anonymous process, tsk->mm == NULL, and tsk->active_mm is the + "borrowed" mm while the anonymous process is running. When the + anonymous process gets scheduled away, the borrowed address space is + returned and cleared. + + To support all that, the "struct mm_struct" now has two counters: a + "mm_users" counter that is how many "real address space users" there are, + and a "mm_count" counter that is the number of "lazy" users (ie anonymous + users) plus one if there are any real users. + + Usually there is at least one real user, but it could be that the real + user exited on another CPU while a lazy user was still active, so you do + actually get cases where you have a address space that is _only_ used by + lazy users. That is often a short-lived state, because once that thread + gets scheduled away in favour of a real thread, the "zombie" mm gets + released because "mm_users" becomes zero. + + Also, a new rule is that _nobody_ ever has "init_mm" as a real MM any + more. "init_mm" should be considered just a "lazy context when no other + context is available", and in fact it is mainly used just at bootup when + no real VM has yet been created. So code that used to check + + if (current->mm == &init_mm) + + should generally just do + + if (!current->mm) + + instead (which makes more sense anyway - the test is basically one of "do + we have a user context", and is generally done by the page fault handler + and things like that). + + Anyway, I put a pre-patch-2.3.13-1 on ftp.kernel.org just a moment ago, + because it slightly changes the interfaces to accommodate the alpha (who + would have thought it, but the alpha actually ends up having one of the + ugliest context switch codes - unlike the other architectures where the MM + and register state is separate, the alpha PALcode joins the two, and you + need to switch both together). + + (From http://marc.info/?l=linux-kernel&m=93337278602211&w=2) diff --git a/Documentation/vm/active_mm.txt b/Documentation/vm/active_mm.txt deleted file mode 100644 index dbf45817405f8..0000000000000 --- a/Documentation/vm/active_mm.txt +++ /dev/null @@ -1,83 +0,0 @@ -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_users" becomes zero. - -Also, a new rule is that _nobody_ ever has "init_mm" as a real MM any -more. "init_mm" should be considered just a "lazy context when no other -context is available", and in fact it is mainly used just at bootup when -no real VM has yet been created. So code that used to check - - if (current->mm == &init_mm) - -should generally just do - - if (!current->mm) - -instead (which makes more sense anyway - the test is basically one of "do -we have a user context", and is generally done by the page fault handler -and things like that). - -Anyway, I put a pre-patch-2.3.13-1 on ftp.kernel.org just a moment ago, -because it slightly changes the interfaces to accommodate the alpha (who -would have thought it, but the alpha actually ends up having one of the -ugliest context switch codes - unlike the other architectures where the MM -and register state is separate, the alpha PALcode joins the two, and you -need to switch both together). - -(From http://marc.info/?l=linux-kernel&m=93337278602211&w=2) diff --git a/Documentation/vm/balance b/Documentation/vm/balance.rst index 964595481af68..6a1fadf3e1735 100644 --- a/Documentation/vm/balance +++ b/Documentation/vm/balance.rst @@ -1,3 +1,9 @@ +.. _balance: + +================ +Memory Balancing +================ + Started Jan 2000 by Kanoj Sarcar <kanoj@sgi.com> Memory balancing is needed for !__GFP_ATOMIC and !__GFP_KSWAPD_RECLAIM as @@ -62,11 +68,11 @@ 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 +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, +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 @@ -89,7 +95,8 @@ 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) + 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) + dma pages. (lkd@tantalophile.demon.co.uk) diff --git a/Documentation/vm/cleancache.txt b/Documentation/vm/cleancache.rst index e4b49df7a0484..68cba9131c318 100644 --- a/Documentation/vm/cleancache.txt +++ b/Documentation/vm/cleancache.rst @@ -1,4 +1,11 @@ -MOTIVATION +.. _cleancache: + +========== +Cleancache +========== + +Motivation +========== Cleancache is a new optional feature provided by the VFS layer that potentially dramatically increases page cache effectiveness for @@ -21,9 +28,10 @@ Transcendent memory "drivers" for cleancache are currently implemented in Xen (using hypervisor memory) and zcache (using in-kernel compressed memory) and other implementations are in development. -FAQs are included below. +:ref:`FAQs <faq>` are included below. -IMPLEMENTATION OVERVIEW +Implementation Overview +======================= A cleancache "backend" that provides transcendent memory registers itself to the kernel's cleancache "frontend" by calling cleancache_register_ops, @@ -80,22 +88,33 @@ different Linux threads are simultaneously putting and invalidating a page with the same handle, the results are indeterminate. Callers must lock the page to ensure serial behavior. -CLEANCACHE PERFORMANCE METRICS +Cleancache Performance Metrics +============================== If properly configured, monitoring of cleancache is done via debugfs in -the /sys/kernel/debug/cleancache directory. The effectiveness of cleancache +the `/sys/kernel/debug/cleancache` directory. The effectiveness of cleancache can be measured (across all filesystems) with: -succ_gets - number of gets that were successful -failed_gets - number of gets that failed -puts - number of puts attempted (all "succeed") -invalidates - number of invalidates attempted +``succ_gets`` + number of gets that were successful + +``failed_gets`` + number of gets that failed + +``puts`` + number of puts attempted (all "succeed") + +``invalidates`` + number of invalidates attempted A backend implementation may provide additional metrics. +.. _faq: + FAQ +=== -1) Where's the value? (Andrew Morton) +* Where's the value? (Andrew Morton) Cleancache provides a significant performance benefit to many workloads in many environments with negligible overhead by improving the @@ -137,8 +156,8 @@ device that stores pages of data in a compressed state. And the proposed "RAMster" driver shares RAM across multiple physical systems. -2) Why does cleancache have its sticky fingers so deep inside the - filesystems and VFS? (Andrew Morton and Christoph Hellwig) +* Why does cleancache have its sticky fingers so deep inside the + filesystems and VFS? (Andrew Morton and Christoph Hellwig) The core hooks for cleancache in VFS are in most cases a single line and the minimum set are placed precisely where needed to maintain @@ -168,9 +187,9 @@ filesystems in the future. The total impact of the hooks to existing fs and mm files is only about 40 lines added (not counting comments and blank lines). -3) Why not make cleancache asynchronous and batched so it can - more easily interface with real devices with DMA instead - of copying each individual page? (Minchan Kim) +* Why not make cleancache asynchronous and batched so it can more + easily interface with real devices with DMA instead of copying each + individual page? (Minchan Kim) The one-page-at-a-time copy semantics simplifies the implementation on both the frontend and backend and also allows the backend to @@ -182,8 +201,8 @@ are avoided. While the interface seems odd for a "real device" or for real kernel-addressable RAM, it makes perfect sense for transcendent memory. -4) Why is non-shared cleancache "exclusive"? And where is the - page "invalidated" after a "get"? (Minchan Kim) +* Why is non-shared cleancache "exclusive"? And where is the + page "invalidated" after a "get"? (Minchan Kim) The main reason is to free up space in transcendent memory and to avoid unnecessary cleancache_invalidate calls. If you want inclusive, @@ -193,7 +212,7 @@ be easily extended to add a "get_no_invalidate" call. The invalidate is done by the cleancache backend implementation. -5) What's the performance impact? +* What's the performance impact? Performance analysis has been presented at OLS'09 and LCA'10. Briefly, performance gains can be significant on most workloads, @@ -206,7 +225,7 @@ single-core systems with slow memory-copy speeds, cleancache has little value, but in newer multicore machines, especially consolidated/virtualized machines, it has great value. -6) How do I add cleancache support for filesystem X? (Boaz Harrash) +* How do I add cleancache support for filesystem X? (Boaz Harrash) Filesystems that are well-behaved and conform to certain restrictions can utilize cleancache simply by making a call to @@ -217,26 +236,26 @@ not enable the optional cleancache. Some points for a filesystem to consider: -- The FS should be block-device-based (e.g. a ram-based FS such - as tmpfs should not enable cleancache) -- To ensure coherency/correctness, the FS must ensure that all - file removal or truncation operations either go through VFS or - add hooks to do the equivalent cleancache "invalidate" operations -- To ensure coherency/correctness, either inode numbers must - be unique across the lifetime of the on-disk file OR the - FS must provide an "encode_fh" function. -- The FS must call the VFS superblock alloc and deactivate routines - or add hooks to do the equivalent cleancache calls done there. -- To maximize performance, all pages fetched from the FS should - go through the do_mpag_readpage routine or the FS should add - hooks to do the equivalent (cf. btrfs) -- Currently, the FS blocksize must be the same as PAGESIZE. This - is not an architectural restriction, but no backends currently - support anything different. -- A clustered FS should invoke the "shared_init_fs" cleancache - hook to get best performance for some backends. - -7) Why not use the KVA of the inode as the key? (Christoph Hellwig) + - The FS should be block-device-based (e.g. a ram-based FS such + as tmpfs should not enable cleancache) + - To ensure coherency/correctness, the FS must ensure that all + file removal or truncation operations either go through VFS or + add hooks to do the equivalent cleancache "invalidate" operations + - To ensure coherency/correctness, either inode numbers must + be unique across the lifetime of the on-disk file OR the + FS must provide an "encode_fh" function. + - The FS must call the VFS superblock alloc and deactivate routines + or add hooks to do the equivalent cleancache calls done there. + - To maximize performance, all pages fetched from the FS should + go through the do_mpag_readpage routine or the FS should add + hooks to do the equivalent (cf. btrfs) + - Currently, the FS blocksize must be the same as PAGESIZE. This + is not an architectural restriction, but no backends currently + support anything different. + - A clustered FS should invoke the "shared_init_fs" cleancache + hook to get best performance for some backends. + +* Why not use the KVA of the inode as the key? (Christoph Hellwig) If cleancache would use the inode virtual address instead of inode/filehandle, the pool id could be eliminated. But, this @@ -251,7 +270,7 @@ of cleancache would be lost because the cache of pages in cleanache is potentially much larger than the kernel pagecache and is most useful if the pages survive inode cache removal. -8) Why is a global variable required? +* Why is a global variable required? The cleancache_enabled flag is checked in all of the frequently-used cleancache hooks. The alternative is a function call to check a static @@ -262,14 +281,14 @@ global variable allows cleancache to be enabled by default at compile time, but have insignificant performance impact when cleancache remains disabled at runtime. -9) Does cleanache work with KVM? +* Does cleanache work with KVM? The memory model of KVM is sufficiently different that a cleancache backend may have less value for KVM. This remains to be tested, especially in an overcommitted system. -10) Does cleancache work in userspace? It sounds useful for - memory hungry caches like web browsers. (Jamie Lokier) +* Does cleancache work in userspace? It sounds useful for + memory hungry caches like web browsers. (Jamie Lokier) No plans yet, though we agree it sounds useful, at least for apps that bypass the page cache (e.g. O_DIRECT). diff --git a/Documentation/vm/conf.py b/Documentation/vm/conf.py new file mode 100644 index 0000000000000..3b0b601af558b --- /dev/null +++ b/Documentation/vm/conf.py @@ -0,0 +1,10 @@ +# -*- coding: utf-8; mode: python -*- + +project = "Linux Memory Management Documentation" + +tags.add("subproject") + +latex_documents = [ + ('index', 'memory-management.tex', project, + 'The kernel development community', 'manual'), +] diff --git a/Documentation/vm/frontswap.txt b/Documentation/vm/frontswap.rst index c71a019be600b..1979f430c1c5b 100644 --- a/Documentation/vm/frontswap.txt +++ b/Documentation/vm/frontswap.rst @@ -1,13 +1,20 @@ +.. _frontswap: + +========= +Frontswap +========= + Frontswap provides a "transcendent memory" interface for swap pages. In some environments, dramatic performance savings may be obtained because swapped pages are saved in RAM (or a RAM-like device) instead of a swap disk. -(Note, frontswap -- and cleancache (merged at 3.0) -- are the "frontends" +(Note, frontswap -- and :ref:`cleancache` (merged at 3.0) -- are the "frontends" and the only necessary changes to the core kernel for transcendent memory; all other supporting code -- the "backends" -- is implemented as drivers. -See the LWN.net article "Transcendent memory in a nutshell" for a detailed -overview of frontswap and related kernel parts: -https://lwn.net/Articles/454795/ ) +See the LWN.net article `Transcendent memory in a nutshell`_ +for a detailed overview of frontswap and related kernel parts) + +.. _Transcendent memory in a nutshell: https://lwn.net/Articles/454795/ Frontswap is so named because it can be thought of as the opposite of a "backing" store for a swap device. The storage is assumed to be @@ -50,19 +57,27 @@ 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 +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 +``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 +=== -1) Where's the value? +* Where's the value? When a workload starts swapping, performance falls through the floor. Frontswap significantly increases performance in many such workloads by @@ -117,8 +132,8 @@ 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. -2) Sure there may be performance advantages in some situations, but - what's the space/time overhead of frontswap? +* 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 @@ -148,8 +163,8 @@ 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. -3) OK, how about a quick overview of what this frontswap patch does - in terms that a kernel hacker can grok? +* 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 @@ -188,9 +203,9 @@ and (potentially) a swap device write are replaced by a "frontswap backend store" and (possibly) a "frontswap backend loads", which are presumably much faster. -4) 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)? +* 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, @@ -240,9 +255,9 @@ 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. -5) Why this weird definition about "duplicate stores"? If a page - has been previously successfully stored, can't it always be - successfully overwritten? +* 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 @@ -254,7 +269,7 @@ 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. -6) What is frontswap_shrink for? +* What is frontswap_shrink for? When the (non-frontswap) swap subsystem swaps out a page to a real swap device, that page is only taking up low-value pre-allocated disk @@ -267,7 +282,7 @@ to "repatriate" pages sent to a remote machine back to the local machine; this is driven using the frontswap_shrink mechanism when memory pressure subsides. -7) Why does the frontswap patch create the new include file swapfile.h? +* 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 diff --git a/Documentation/vm/highmem.txt b/Documentation/vm/highmem.rst index 4324d24ffacd1..0f69a9fec34dd 100644 --- a/Documentation/vm/highmem.txt +++ b/Documentation/vm/highmem.rst @@ -1,25 +1,14 @@ +.. _highmem: - ==================== - HIGH MEMORY HANDLING - ==================== +==================== +High Memory Handling +==================== By: Peter Zijlstra <a.p.zijlstra@chello.nl> -Contents: - - (*) What is high memory? - - (*) Temporary virtual mappings. - - (*) Using kmap_atomic. - - (*) Cost of temporary mappings. - - (*) i386 PAE. +.. contents:: :local: - -==================== -WHAT IS HIGH MEMORY? +What Is High Memory? ==================== High memory (highmem) is used when the size of physical memory approaches or @@ -38,7 +27,7 @@ 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: +userspace and the top 1GiB for kernel space:: +--------+ 0xffffffff | Kernel | @@ -58,40 +47,38 @@ and user maps. Some hardware (like some ARMs), however, have limited virtual space when they use mm context tags. -========================== -TEMPORARY VIRTUAL MAPPINGS +Temporary Virtual Mappings ========================== The kernel contains several ways of creating temporary mappings: - (*) vmap(). This can be used to make a long duration mapping of multiple - physical pages into a contiguous virtual space. It needs global - synchronization to unmap. +* vmap(). This can be used to make a long duration mapping of multiple + physical pages into a contiguous virtual space. It needs global + synchronization to unmap. - (*) kmap(). This permits a short duration mapping of a single page. It needs - global synchronization, but is amortized somewhat. It is also prone to - deadlocks when using in a nested fashion, and so it is not recommended for - new code. +* kmap(). This permits a short duration mapping of a single page. It needs + global synchronization, but is amortized somewhat. It is also prone to + deadlocks when using in a nested fashion, and so it is not recommended for + new code. - (*) kmap_atomic(). This permits a very short duration mapping of a single - page. Since the mapping is restricted to the CPU that issued it, it - performs well, but the issuing task is therefore required to stay on that - CPU until it has finished, lest some other task displace its mappings. +* kmap_atomic(). This permits a very short duration mapping of a single + page. Since the mapping is restricted to the CPU that issued it, it + performs well, but the issuing task is therefore required to stay on that + CPU until it has finished, lest some other task displace its mappings. - kmap_atomic() may also be used by interrupt contexts, since it is does not - sleep and the caller may not sleep until after kunmap_atomic() is called. + kmap_atomic() may also be used by interrupt contexts, since it is does not + sleep and the caller may not sleep until after kunmap_atomic() is called. - It may be assumed that k[un]map_atomic() won't fail. + It may be assumed that k[un]map_atomic() won't fail. -================= -USING KMAP_ATOMIC +Using kmap_atomic ================= When and where to use kmap_atomic() is straightforward. It is used when code wants to access the contents of a page that might be allocated from high memory (see __GFP_HIGHMEM), for example a page in the pagecache. The API has two -functions, and they can be used in a manner similar to the following: +functions, and they can be used in a manner similar to the following:: /* Find the page of interest. */ struct page *page = find_get_page(mapping, offset); @@ -109,7 +96,7 @@ Note that the kunmap_atomic() call takes the result of the kmap_atomic() call not the argument. If you need to map two pages because you want to copy from one page to -another you need to keep the kmap_atomic calls strictly nested, like: +another you need to keep the kmap_atomic calls strictly nested, like:: vaddr1 = kmap_atomic(page1); vaddr2 = kmap_atomic(page2); @@ -120,8 +107,7 @@ another you need to keep the kmap_atomic calls strictly nested, like: kunmap_atomic(vaddr1); -========================== -COST OF TEMPORARY MAPPINGS +Cost of Temporary Mappings ========================== The cost of creating temporary mappings can be quite high. The arch has to @@ -136,25 +122,24 @@ 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: +* 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... +* 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. +* 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 diff --git a/Documentation/vm/hmm.txt b/Documentation/vm/hmm.rst index 2d1d6f69e91bd..cdf3911582c86 100644 --- a/Documentation/vm/hmm.txt +++ b/Documentation/vm/hmm.rst @@ -1,4 +1,8 @@ +.. 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 @@ -6,10 +10,10 @@ 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 address 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 hetero- -geneous computing where GPU, DSP, or FPGA are used to perform various +allowing a device to transparently access program address 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 @@ -21,19 +25,10 @@ fifth section deals with how device memory is represented inside the kernel. Finally, the last section presents a new migration helper that allows lever- aging the device DMA engine. +.. contents:: :local: -1) Problems of using a device specific memory allocator: -2) I/O bus, device memory characteristics -3) Shared address space and migration -4) Address space mirroring implementation and API -5) Represent and manage device memory from core kernel point of view -6) Migration to and from device memory -7) Memory cgroup (memcg) and rss accounting - - -------------------------------------------------------------------------------- - -1) Problems of using a device specific memory allocator: +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. @@ -77,9 +72,8 @@ are only do-able with a shared address space. It is also more reasonable to use a shared address space for all other patterns. -------------------------------------------------------------------------------- - -2) I/O bus, device memory characteristics +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 @@ -109,9 +103,8 @@ access any memory but we must also permit any memory to be migrated to device memory while device is using it (blocking CPU access while it happens). -------------------------------------------------------------------------------- - -3) Shared address space and migration +Shared address space and migration +================================== HMM intends to provide two main features. First one is to share the address space by duplicating the CPU page table in the device page table so the same @@ -148,23 +141,23 @@ ages device memory by migrating the part of the data set that is actively being used by the device. -------------------------------------------------------------------------------- - -4) Address space mirroring implementation and API +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 an hmm_mirror struct: +registration of an hmm_mirror struct:: int hmm_mirror_register(struct hmm_mirror *mirror, struct mm_struct *mm); int hmm_mirror_register_locked(struct hmm_mirror *mirror, struct mm_struct *mm); + The locked variant is to be used when the driver is already holding mmap_sem of the mm in write mode. The mirror struct has a set of callbacks that are used -to propagate CPU page tables: +to propagate CPU page tables:: struct hmm_mirror_ops { /* sync_cpu_device_pagetables() - synchronize page tables @@ -193,10 +186,10 @@ The device driver must perform the update action to the range (mark range read only, or fully unmap, ...). The device must be done with the update before the driver callback returns. - When the device driver wants to populate a range of virtual addresses, it can -use either: - int hmm_vma_get_pfns(struct vm_area_struct *vma, +use either:: + + int hmm_vma_get_pfns(struct vm_area_struct *vma, struct hmm_range *range, unsigned long start, unsigned long end, @@ -221,7 +214,7 @@ provides a set of flags to help the driver identify special CPU page table entries. Locking with the update() callback is the most important aspect the driver must -respect in order to keep things properly synchronized. The usage pattern is: +respect in order to keep things properly synchronized. The usage pattern is:: int driver_populate_range(...) { @@ -262,9 +255,8 @@ report commands as executed is serialized (there is no point in doing this concurrently). -------------------------------------------------------------------------------- - -5) Represent and manage device memory from core kernel point of view +Represent and manage device memory from core kernel point of view +================================================================= Several different designs were tried to support device memory. First one used a device specific data structure to keep information about migrated memory and @@ -280,14 +272,14 @@ unaware of the difference. We only need to make sure that no one ever tries to map those pages from the CPU side. HMM provides a set of helpers to register and hotplug device memory as a new -region needing a struct page. This is offered through a very simple API: +region needing a struct page. This is offered through a very simple API:: struct hmm_devmem *hmm_devmem_add(const struct hmm_devmem_ops *ops, struct device *device, unsigned long size); void hmm_devmem_remove(struct hmm_devmem *devmem); -The hmm_devmem_ops is where most of the important things are: +The hmm_devmem_ops is where most of the important things are:: struct hmm_devmem_ops { void (*free)(struct hmm_devmem *devmem, struct page *page); @@ -306,13 +298,12 @@ which it cannot do. This second callback must trigger a migration back to system memory. -------------------------------------------------------------------------------- - -6) Migration to and from device memory +Migration to and from device memory +=================================== Because the CPU cannot access device memory, migration must use the device DMA engine to perform copy from and to device memory. For this we need a new -migration helper: +migration helper:: int migrate_vma(const struct migrate_vma_ops *ops, struct vm_area_struct *vma, @@ -331,7 +322,7 @@ migration might be for a range of addresses the device is actively accessing. The migrate_vma_ops struct defines two callbacks. First one (alloc_and_copy()) controls destination memory allocation and copy operation. Second one is there -to allow the device driver to perform cleanup operations after migration. +to allow the device driver to perform cleanup operations after migration:: struct migrate_vma_ops { void (*alloc_and_copy)(struct vm_area_struct *vma, @@ -365,9 +356,8 @@ bandwidth but this is considered as a rare event and a price that we are willing to pay to keep all the code simpler. -------------------------------------------------------------------------------- - -7) Memory cgroup (memcg) and rss accounting +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 diff --git a/Documentation/vm/hugetlbfs_reserv.txt b/Documentation/vm/hugetlbfs_reserv.rst index 9aca09a76bed5..9d200762114f6 100644 --- a/Documentation/vm/hugetlbfs_reserv.txt +++ b/Documentation/vm/hugetlbfs_reserv.rst @@ -1,6 +1,13 @@ -Hugetlbfs Reservation Overview ------------------------------- -Huge pages as described at 'Documentation/vm/hugetlbpage.txt' are typically +.. _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 @@ -17,47 +24,55 @@ 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). + 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; - }; + 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 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. + + ``#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 @@ -65,12 +80,14 @@ Page Flags 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 two types of mappings. Location differences are: + - For private mappings, the reservation map hangs off the the VMA structure. Specifically, vma->vm_private_data. This reserve map is created at the time the mapping (mmap(MAP_PRIVATE)) is created. @@ -82,15 +99,15 @@ of mappings. Location differences are: 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() +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) + 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 for the NORESERVE flag was specified in either the shmget() or mmap() call. If NORESERVE @@ -105,6 +122,7 @@ 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. @@ -121,12 +139,13 @@ 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 "Reservation Map Modifications" for details on how this is accomplished. +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 "Subpool Reservations" for more details. +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 @@ -135,9 +154,11 @@ 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. +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. @@ -152,14 +173,18 @@ 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) +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 @@ -170,8 +195,9 @@ 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 -"Reservation Map Helper Routines" for detailed information on what this -routine does. The value returned from vma_needs_reservation() is generally +: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. @@ -180,21 +206,25 @@ 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: +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 @@ -222,12 +252,14 @@ 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 "Reservation Map Helper Routines" for more +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 @@ -237,7 +269,8 @@ 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 @@ -247,7 +280,8 @@ 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 "Consuming Reservations/Allocating a Huge Page" +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 @@ -259,9 +293,11 @@ 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 @@ -295,7 +331,8 @@ 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 @@ -326,30 +363,36 @@ 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); +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. -2a) 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(). -2b) If the operation can not succeed, region_abort is called for the same range - [f, t) to abort the operation. +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 @@ -371,6 +414,7 @@ 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 @@ -384,6 +428,7 @@ It is typically called in the following situations: 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 @@ -403,9 +448,11 @@ 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, @@ -414,32 +461,40 @@ 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. +shared mappings and hide this detail from the caller:: + + long vma_needs_reservation(struct hstate *h, + struct vm_area_struct *vma, + unsigned long addr) -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. +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) -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. +map was modified between calls:: + + void vma_end_reservation(struct hstate *h, + struct vm_area_struct *vma, + unsigned long addr) -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. +operation:: + + long vma_add_reservation(struct hstate *h, + struct vm_area_struct *vma, + unsigned long addr) -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 @@ -453,8 +508,10 @@ be done on error paths. Reservation Cleanup in Error Paths ----------------------------------- -As mentioned in the section "Reservation Map Helper Routines", reservation +================================== + +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. @@ -494,13 +551,14 @@ 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. +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 @@ -525,5 +583,13 @@ of cpusets or memory policy there is no guarantee that huge pages will be available on the required nodes. This is true even if there are a sufficient number of global reservations. +Hugetlbfs regression testing +============================ + +The most complete set of hugetlb tests are in the libhugetlbfs repository. +If you modify any hugetlb related code, use the libhugetlbfs test suite +to check for regressions. In addition, if you add any new hugetlb +functionality, please add appropriate tests to libhugetlbfs. +-- Mike Kravetz, 7 April 2017 diff --git a/Documentation/vm/hugetlbpage.txt b/Documentation/vm/hugetlbpage.txt deleted file mode 100644 index faf077d50d427..0000000000000 --- a/Documentation/vm/hugetlbpage.txt +++ /dev/null @@ -1,351 +0,0 @@ - -The intent of this file is to give a brief summary of hugetlbpage support in -the Linux kernel. This support is built on top of multiple page size support -that is provided by most modern architectures. For example, x86 CPUs normally -support 4K and 2M (1G if architecturally supported) page sizes, ia64 -architecture supports multiple page sizes 4K, 8K, 64K, 256K, 1M, 4M, 16M, -256M and ppc64 supports 4K and 16M. A TLB is a cache of virtual-to-physical -translations. Typically this is a very scarce resource on processor. -Operating systems try to make best use of limited number of TLB resources. -This optimization is more critical now as bigger and bigger physical memories -(several GBs) are more readily available. - -Users can use the huge page support in Linux kernel by either using the mmap -system call or standard SYSV shared memory system calls (shmget, shmat). - -First the Linux kernel needs to be built with the CONFIG_HUGETLBFS -(present under "File systems") and CONFIG_HUGETLB_PAGE (selected -automatically when CONFIG_HUGETLBFS is selected) configuration -options. - -The /proc/meminfo file provides information about the total number of -persistent hugetlb pages in the kernel's huge page pool. It also displays -default huge page size and information about the number of free, reserved -and surplus huge pages in the pool of huge pages of default size. -The huge page size is needed for generating the proper alignment and -size of the arguments to system calls that map huge page regions. - -The output of "cat /proc/meminfo" will include lines like: - -..... -HugePages_Total: uuu -HugePages_Free: vvv -HugePages_Rsvd: www -HugePages_Surp: xxx -Hugepagesize: yyy kB -Hugetlb: zzz kB - -where: -HugePages_Total is the size of the pool of huge pages. -HugePages_Free is the number of huge pages in the pool that are not yet - allocated. -HugePages_Rsvd is short for "reserved," and is the number of huge pages for - which a commitment to allocate from the pool has been made, - but no allocation has yet been made. Reserved huge pages - guarantee that an application will be able to allocate a - huge page from the pool of huge pages at fault time. -HugePages_Surp is short for "surplus," and is the number of huge pages in - the pool above the value in /proc/sys/vm/nr_hugepages. The - maximum number of surplus huge pages is controlled by - /proc/sys/vm/nr_overcommit_hugepages. -Hugepagesize is the default hugepage size (in Kb). -Hugetlb is the total amount of memory (in kB), consumed by huge - pages of all sizes. - If huge pages of different sizes are in use, this number - will exceed HugePages_Total * Hugepagesize. To get more - detailed information, please, refer to - /sys/kernel/mm/hugepages (described below). - - -/proc/filesystems should also show a filesystem of type "hugetlbfs" configured -in the kernel. - -/proc/sys/vm/nr_hugepages indicates the current number of "persistent" huge -pages in the kernel's huge page pool. "Persistent" huge pages will be -returned to the huge page pool when freed by a task. A user with root -privileges can dynamically allocate more or free some persistent huge pages -by increasing or decreasing the value of 'nr_hugepages'. - -Pages that are used as huge pages are reserved inside the kernel and cannot -be used for other purposes. Huge pages cannot be swapped out under -memory pressure. - -Once a number of huge pages have been pre-allocated to the kernel huge page -pool, a user with appropriate privilege can use either the mmap system call -or shared memory system calls to use the huge pages. See the discussion of -Using Huge Pages, below. - -The administrator can allocate persistent huge pages on the kernel boot -command line by specifying the "hugepages=N" parameter, where 'N' = the -number of huge pages requested. This is the most reliable method of -allocating huge pages as memory has not yet become fragmented. - -Some platforms support multiple huge page sizes. To allocate huge pages -of a specific size, one must precede the huge pages boot command parameters -with a huge page size selection parameter "hugepagesz=<size>". <size> must -be specified in bytes with optional scale suffix [kKmMgG]. The default huge -page size may be selected with the "default_hugepagesz=<size>" boot parameter. - -When multiple huge page sizes are supported, /proc/sys/vm/nr_hugepages -indicates the current number of pre-allocated huge pages of the default size. -Thus, one can use the following command to dynamically allocate/deallocate -default sized persistent huge pages: - - echo 20 > /proc/sys/vm/nr_hugepages - -This command will try to adjust the number of default sized huge pages in the -huge page pool to 20, allocating or freeing huge pages, as required. - -On a NUMA platform, the kernel will attempt to distribute the huge page pool -over all the set of allowed nodes specified by the NUMA memory policy of the -task that modifies nr_hugepages. The default for the allowed nodes--when the -task has default memory policy--is all on-line nodes with memory. Allowed -nodes with insufficient available, contiguous memory for a huge page will be -silently skipped when allocating persistent huge pages. See the discussion -below of the interaction of task memory policy, cpusets and per node attributes -with the allocation and freeing of persistent huge pages. - -The success or failure of huge page allocation depends on the amount of -physically contiguous memory that is present in system at the time of the -allocation attempt. If the kernel is unable to allocate huge pages from -some nodes in a NUMA system, it will attempt to make up the difference by -allocating extra pages on other nodes with sufficient available contiguous -memory, if any. - -System administrators may want to put this command in one of the local rc -init files. This will enable the kernel to allocate huge pages early in -the boot process when the possibility of getting physical contiguous pages -is still very high. Administrators can verify the number of huge pages -actually allocated by checking the sysctl or meminfo. To check the per node -distribution of huge pages in a NUMA system, use: - - cat /sys/devices/system/node/node*/meminfo | fgrep Huge - -/proc/sys/vm/nr_overcommit_hugepages specifies how large the pool of -huge pages can grow, if more huge pages than /proc/sys/vm/nr_hugepages are -requested by applications. Writing any non-zero value into this file -indicates that the hugetlb subsystem is allowed to try to obtain that -number of "surplus" huge pages from the kernel's normal page pool, when the -persistent huge page pool is exhausted. As these surplus huge pages become -unused, they are freed back to the kernel's normal page pool. - -When increasing the huge page pool size via nr_hugepages, any existing surplus -pages will first be promoted to persistent huge pages. Then, additional -huge pages will be allocated, if necessary and if possible, to fulfill -the new persistent huge page pool size. - -The administrator may shrink the pool of persistent huge pages for -the default huge page size by setting the nr_hugepages sysctl to a -smaller value. The kernel will attempt to balance the freeing of huge pages -across all nodes in the memory policy of the task modifying nr_hugepages. -Any free huge pages on the selected nodes will be freed back to the kernel's -normal page pool. - -Caveat: Shrinking the persistent huge page pool via nr_hugepages such that -it becomes less than the number of huge pages in use will convert the balance -of the in-use huge pages to surplus huge pages. This will occur even if -the number of surplus pages it would exceed the overcommit value. As long as -this condition holds--that is, until nr_hugepages+nr_overcommit_hugepages is -increased sufficiently, or the surplus huge pages go out of use and are freed-- -no more surplus huge pages will be allowed to be allocated. - -With support for multiple huge page pools at run-time available, much of -the huge page userspace interface in /proc/sys/vm has been duplicated in sysfs. -The /proc interfaces discussed above have been retained for backwards -compatibility. The root huge page control directory in sysfs is: - - /sys/kernel/mm/hugepages - -For each huge page size supported by the running kernel, a subdirectory -will exist, of the form: - - hugepages-${size}kB - -Inside each of these directories, the same set of files will exist: - - nr_hugepages - nr_hugepages_mempolicy - nr_overcommit_hugepages - free_hugepages - resv_hugepages - surplus_hugepages - -which function as described above for the default huge page-sized case. - - -Interaction of Task Memory Policy with Huge Page Allocation/Freeing -=================================================================== - -Whether huge pages are allocated and freed via the /proc interface or -the /sysfs interface using the nr_hugepages_mempolicy attribute, the NUMA -nodes from which huge pages are allocated or freed are controlled by the -NUMA memory policy of the task that modifies the nr_hugepages_mempolicy -sysctl or attribute. When the nr_hugepages attribute is used, mempolicy -is ignored. - -The recommended method to allocate or free huge pages to/from the kernel -huge page pool, using the nr_hugepages example above, is: - - numactl --interleave <node-list> echo 20 \ - >/proc/sys/vm/nr_hugepages_mempolicy - -or, more succinctly: - - numactl -m <node-list> echo 20 >/proc/sys/vm/nr_hugepages_mempolicy - -This will allocate or free abs(20 - nr_hugepages) to or from the nodes -specified in <node-list>, depending on whether number of persistent huge pages -is initially less than or greater than 20, respectively. No huge pages will be -allocated nor freed on any node not included in the specified <node-list>. - -When adjusting the persistent hugepage count via nr_hugepages_mempolicy, any -memory policy mode--bind, preferred, local or interleave--may be used. The -resulting effect on persistent huge page allocation is as follows: - -1) Regardless of mempolicy mode [see Documentation/vm/numa_memory_policy.txt], - persistent huge pages will be distributed across the node or nodes - specified in the mempolicy as if "interleave" had been specified. - However, if a node in the policy does not contain sufficient contiguous - memory for a huge page, the allocation will not "fallback" to the nearest - neighbor node with sufficient contiguous memory. To do this would cause - undesirable imbalance in the distribution of the huge page pool, or - possibly, allocation of persistent huge pages on nodes not allowed by - the task's memory policy. - -2) One or more nodes may be specified with the bind or interleave policy. - If more than one node is specified with the preferred policy, only the - lowest numeric id will be used. Local policy will select the node where - the task is running at the time the nodes_allowed mask is constructed. - For local policy to be deterministic, the task must be bound to a cpu or - cpus in a single node. Otherwise, the task could be migrated to some - other node at any time after launch and the resulting node will be - indeterminate. Thus, local policy is not very useful for this purpose. - Any of the other mempolicy modes may be used to specify a single node. - -3) The nodes allowed mask will be derived from any non-default task mempolicy, - whether this policy was set explicitly by the task itself or one of its - ancestors, such as numactl. This means that if the task is invoked from a - shell with non-default policy, that policy will be used. One can specify a - node list of "all" with numactl --interleave or --membind [-m] to achieve - interleaving over all nodes in the system or cpuset. - -4) Any task mempolicy specified--e.g., using numactl--will be constrained by - the resource limits of any cpuset in which the task runs. Thus, there will - be no way for a task with non-default policy running in a cpuset with a - subset of the system nodes to allocate huge pages outside the cpuset - without first moving to a cpuset that contains all of the desired nodes. - -5) Boot-time huge page allocation attempts to distribute the requested number - of huge pages over all on-lines nodes with memory. - -Per Node Hugepages Attributes -============================= - -A subset of the contents of the root huge page control directory in sysfs, -described above, will be replicated under each the system device of each -NUMA node with memory in: - - /sys/devices/system/node/node[0-9]*/hugepages/ - -Under this directory, the subdirectory for each supported huge page size -contains the following attribute files: - - nr_hugepages - free_hugepages - surplus_hugepages - -The free_' and surplus_' attribute files are read-only. They return the number -of free and surplus [overcommitted] huge pages, respectively, on the parent -node. - -The nr_hugepages attribute returns the total number of huge pages on the -specified node. When this attribute is written, the number of persistent huge -pages on the parent node will be adjusted to the specified value, if sufficient -resources exist, regardless of the task's mempolicy or cpuset constraints. - -Note that the number of overcommit and reserve pages remain global quantities, -as we don't know until fault time, when the faulting task's mempolicy is -applied, from which node the huge page allocation will be attempted. - - -Using Huge Pages -================ - -If the user applications are going to request huge pages using mmap system -call, then it is required that system administrator mount a file system of -type hugetlbfs: - - mount -t hugetlbfs \ - -o uid=<value>,gid=<value>,mode=<value>,pagesize=<value>,size=<value>,\ - min_size=<value>,nr_inodes=<value> none /mnt/huge - -This command mounts a (pseudo) filesystem of type hugetlbfs on the directory -/mnt/huge. Any files created on /mnt/huge uses huge pages. The uid and gid -options sets the owner and group of the root of the file system. By default -the uid and gid of the current process are taken. The mode option sets the -mode of root of file system to value & 01777. This value is given in octal. -By default the value 0755 is picked. If the platform supports multiple huge -page sizes, the pagesize option can be used to specify the huge page size and -associated pool. pagesize is specified in bytes. If pagesize is not specified -the platform's default huge page size and associated pool will be used. The -size option sets the maximum value of memory (huge pages) allowed for that -filesystem (/mnt/huge). The size option can be specified in bytes, or as a -percentage of the specified huge page pool (nr_hugepages). The size is -rounded down to HPAGE_SIZE boundary. The min_size option sets the minimum -value of memory (huge pages) allowed for the filesystem. min_size can be -specified in the same way as size, either bytes or a percentage of the -huge page pool. At mount time, the number of huge pages specified by -min_size are reserved for use by the filesystem. If there are not enough -free huge pages available, the mount will fail. As huge pages are allocated -to the filesystem and freed, the reserve count is adjusted so that the sum -of allocated and reserved huge pages is always at least min_size. The option -nr_inodes sets the maximum number of inodes that /mnt/huge can use. If the -size, min_size or nr_inodes option is not provided on command line then -no limits are set. For pagesize, size, min_size and nr_inodes options, you -can use [G|g]/[M|m]/[K|k] to represent giga/mega/kilo. For example, size=2K -has the same meaning as size=2048. - -While read system calls are supported on files that reside on hugetlb -file systems, write system calls are not. - -Regular chown, chgrp, and chmod commands (with right permissions) could be -used to change the file attributes on hugetlbfs. - -Also, it is important to note that no such mount command is required if -applications are going to use only shmat/shmget system calls or mmap with -MAP_HUGETLB. For an example of how to use mmap with MAP_HUGETLB see map_hugetlb -below. - -Users who wish to use hugetlb memory via shared memory segment should be a -member of a supplementary group and system admin needs to configure that gid -into /proc/sys/vm/hugetlb_shm_group. It is possible for same or different -applications to use any combination of mmaps and shm* calls, though the mount of -filesystem will be required for using mmap calls without MAP_HUGETLB. - -Syscalls that operate on memory backed by hugetlb pages only have their lengths -aligned to the native page size of the processor; they will normally fail with -errno set to EINVAL or exclude hugetlb pages that extend beyond the length if -not hugepage aligned. For example, munmap(2) will fail if memory is backed by -a hugetlb page and the length is smaller than the hugepage size. - - -Examples -======== - -1) map_hugetlb: see tools/testing/selftests/vm/map_hugetlb.c - -2) hugepage-shm: see tools/testing/selftests/vm/hugepage-shm.c - -3) hugepage-mmap: see tools/testing/selftests/vm/hugepage-mmap.c - -4) The libhugetlbfs (https://github.com/libhugetlbfs/libhugetlbfs) library - provides a wide range of userspace tools to help with huge page usability, - environment setup, and control. - -Kernel development 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. diff --git a/Documentation/vm/hwpoison.txt b/Documentation/vm/hwpoison.rst index e912d7eee769a..09bd24a927842 100644 --- a/Documentation/vm/hwpoison.txt +++ b/Documentation/vm/hwpoison.rst @@ -1,7 +1,14 @@ +.. 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", +(``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. @@ -46,9 +53,10 @@ 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) modi memory failure recovery can be in: +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. @@ -67,9 +75,8 @@ late kill This is best for memory error unaware applications and default Note some pages are always handled as late kill. ---- - -User control: +User control +============ vm.memory_failure_recovery See sysctl.txt @@ -79,11 +86,19 @@ vm.memory_failure_early_kill 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 + + 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, @@ -92,77 +107,64 @@ PR_MCE_KILL PR_MCE_KILL_GET return current mode +Testing +======= ---- - -Testing: - -madvise(MADV_HWPOISON, ....) - (as root) - Poison a page in the process for testing - +* madvise(MADV_HWPOISON, ....) (as root) - Poison a page in the + process for testing -hwpoison-inject module through debugfs +* hwpoison-inject module through debugfs ``/sys/kernel/debug/hwpoison/`` -/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. -corrupt-pfn + 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. -Inject hwpoison fault at PFN echoed into this file. This does -some early filtering to avoid corrupted unintended pages in test suites. + Note these injection interfaces are not stable and might change between + kernel versions -unpoison-pfn + 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. -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. + corrupt-filter-memcg + Limit injection to pages owned by memgroup. Specified by inode + number of the memcg. -Note these injection interfaces are not stable and might change between -kernel versions + Example:: -corrupt-filter-dev-major -corrupt-filter-dev-minor + mkdir /sys/fs/cgroup/mem/hwpoison -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. + usemem -m 100 -s 1000 & + echo `jobs -p` > /sys/fs/cgroup/mem/hwpoison/tasks -corrupt-filter-memcg + memcg_ino=$(ls -id /sys/fs/cgroup/mem/hwpoison | cut -f1 -d' ') + echo $memcg_ino > /debug/hwpoison/corrupt-filter-memcg -Limit injection to pages owned by memgroup. Specified by inode number -of the memcg. + page-types -p `pidof init` --hwpoison # shall do nothing + page-types -p `pidof usemem` --hwpoison # poison its pages -Example: - mkdir /sys/fs/cgroup/mem/hwpoison + corrupt-filter-flags-mask, corrupt-filter-flags-value + When specified, only poison pages if ((page_flags & mask) == + value). This allows stress testing of many kinds of + pages. The page_flags are the same as in /proc/kpageflags. The + flag bits are defined in include/linux/kernel-page-flags.h and + documented in Documentation/admin-guide/mm/pagemap.rst - usemem -m 100 -s 1000 & - echo `jobs -p` > /sys/fs/cgroup/mem/hwpoison/tasks +* Architecture specific MCE injector - memcg_ino=$(ls -id /sys/fs/cgroup/mem/hwpoison | cut -f1 -d' ') - echo $memcg_ino > /debug/hwpoison/corrupt-filter-memcg + x86 has mce-inject, mce-test - page-types -p `pidof init` --hwpoison # shall do nothing - page-types -p `pidof usemem` --hwpoison # poison its pages + Some portable hwpoison test programs in mce-test, see below. -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/vm/pagemap.txt - -Architecture specific MCE injector - -x86 has mce-inject, mce-test - -Some portable hwpoison test programs in mce-test, see blow. - ---- - -References: +References +========== http://halobates.de/mce-lc09-2.pdf Overview presentation from LinuxCon 09 @@ -174,14 +176,11 @@ git://git.kernel.org/pub/scm/utils/cpu/mce/mce-inject.git x86 specific injector ---- - -Limitations: - +Limitations +=========== - Not all page types are supported and never will. Most kernel internal -objects cannot be recovered, only LRU pages for now. + objects cannot be recovered, only LRU pages for now. - Right now hugepage support is missing. --- Andi Kleen, Oct 2009 - diff --git a/Documentation/vm/idle_page_tracking.txt b/Documentation/vm/idle_page_tracking.txt deleted file mode 100644 index 85dcc3bb85dca..0000000000000 --- a/Documentation/vm/idle_page_tracking.txt +++ /dev/null @@ -1,98 +0,0 @@ -MOTIVATION - -The idle page tracking feature allows to track which memory pages are being -accessed by a workload and which are idle. This information can be useful for -estimating the workload's working set size, which, in turn, can be taken into -account when configuring the workload parameters, setting memory cgroup limits, -or deciding where to place the workload within a compute cluster. - -It is enabled by CONFIG_IDLE_PAGE_TRACKING=y. - -USER API - -The idle page tracking API is located at /sys/kernel/mm/page_idle. Currently, -it consists of the only read-write file, /sys/kernel/mm/page_idle/bitmap. - -The file implements a bitmap where each bit corresponds to a memory page. The -bitmap is represented by an array of 8-byte integers, and the page at PFN #i is -mapped to bit #i%64 of array element #i/64, byte order is native. When a bit is -set, the corresponding page is idle. - -A page is considered idle if it has not been accessed since it was marked idle -(for more details on what "accessed" actually means see the IMPLEMENTATION -DETAILS section). To mark a page idle one has to set the bit corresponding to -the page by writing to the file. A value written to the file is OR-ed with the -current bitmap value. - -Only accesses to user memory pages are tracked. These are pages mapped to a -process address space, page cache and buffer pages, swap cache pages. For other -page types (e.g. SLAB pages) an attempt to mark a page idle is silently ignored, -and hence such pages are never reported idle. - -For huge pages the idle flag is set only on the head page, so one has to read -/proc/kpageflags in order to correctly count idle huge pages. - -Reading from or writing to /sys/kernel/mm/page_idle/bitmap will return --EINVAL if you are not starting the read/write on an 8-byte boundary, or -if the size of the read/write is not a multiple of 8 bytes. Writing to -this file beyond max PFN will return -ENXIO. - -That said, in order to estimate the amount of pages that are not used by a -workload one should: - - 1. Mark all the workload's pages as idle by setting corresponding bits in - /sys/kernel/mm/page_idle/bitmap. The pages can be found by reading - /proc/pid/pagemap if the workload is represented by a process, or by - filtering out alien pages using /proc/kpagecgroup in case the workload is - placed in a memory cgroup. - - 2. Wait until the workload accesses its working set. - - 3. Read /sys/kernel/mm/page_idle/bitmap and count the number of bits set. If - one wants to ignore certain types of pages, e.g. mlocked pages since they - are not reclaimable, he or she can filter them out using /proc/kpageflags. - -See Documentation/vm/pagemap.txt for more information about /proc/pid/pagemap, -/proc/kpageflags, and /proc/kpagecgroup. - -IMPLEMENTATION DETAILS - -The kernel internally keeps track of accesses to user memory pages in order to -reclaim unreferenced pages first on memory shortage conditions. A page is -considered referenced if it has been recently accessed via a process address -space, in which case one or more PTEs it is mapped to will have the Accessed bit -set, or marked accessed explicitly by the kernel (see mark_page_accessed()). The -latter happens when: - - - a userspace process reads or writes a page using a system call (e.g. read(2) - or write(2)) - - - a page that is used for storing filesystem buffers is read or written, - because a process needs filesystem metadata stored in it (e.g. lists a - directory tree) - - - a page is accessed by a device driver using get_user_pages() - -When a dirty page is written to swap or disk as a result of memory reclaim or -exceeding the dirty memory limit, it is not marked referenced. - -The idle memory tracking feature adds a new page flag, the Idle flag. This flag -is set manually, by writing to /sys/kernel/mm/page_idle/bitmap (see the USER API -section), and cleared automatically whenever a page is referenced as defined -above. - -When a page is marked idle, the Accessed bit must be cleared in all PTEs it is -mapped to, otherwise we will not be able to detect accesses to the page coming -from a process address space. To avoid interference with the reclaimer, which, -as noted above, uses the Accessed bit to promote actively referenced pages, one -more page flag is introduced, the Young flag. When the PTE Accessed bit is -cleared as a result of setting or updating a page's Idle flag, the Young flag -is set on the page. The reclaimer treats the Young flag as an extra PTE -Accessed bit and therefore will consider such a page as referenced. - -Since the idle memory tracking feature is based on the memory reclaimer logic, -it only works with pages that are on an LRU list, other pages are silently -ignored. That means it will ignore a user memory page if it is isolated, but -since there are usually not many of them, it should not affect the overall -result noticeably. In order not to stall scanning of the idle page bitmap, -locked pages may be skipped too. diff --git a/Documentation/vm/index.rst b/Documentation/vm/index.rst new file mode 100644 index 0000000000000..c4ded22197ca1 --- /dev/null +++ b/Documentation/vm/index.rst @@ -0,0 +1,50 @@ +===================================== +Linux Memory Management Documentation +===================================== + +This is a collection of documents about Linux memory management (mm) subsystem. + +User guides for MM features +=========================== + +The following documents provide guides for controlling and tuning +various features of the Linux memory management + +.. toctree:: + :maxdepth: 1 + + swap_numa + zswap + +Kernel developers MM documentation +================================== + +The below documents describe MM internals with different level of +details ranging from notes and mailing list responses to elaborate +descriptions of data structures and algorithms. + +.. toctree:: + :maxdepth: 1 + + active_mm + balance + cleancache + frontswap + highmem + hmm + hwpoison + hugetlbfs_reserv + ksm + mmu_notifier + numa + overcommit-accounting + page_migration + page_frags + page_owner + remap_file_pages + slub + split_page_table_lock + transhuge + unevictable-lru + z3fold + zsmalloc diff --git a/Documentation/vm/ksm.rst b/Documentation/vm/ksm.rst new file mode 100644 index 0000000000000..d32016d9be2cb --- /dev/null +++ b/Documentation/vm/ksm.rst @@ -0,0 +1,87 @@ +.. _ksm: + +======================= +Kernel Samepage Merging +======================= + +KSM is a memory-saving de-duplication feature, enabled by CONFIG_KSM=y, +added to the Linux kernel in 2.6.32. See ``mm/ksm.c`` for its implementation, +and http://lwn.net/Articles/306704/ and http://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 :c:type:`struct rmap_item` and the ``page->mapping`` of the +KSM page points to the stable tree node. + +When the sharing passes this threshold, KSM adds a second dimension to +the stable tree. The tree node becomes a "chain" that links one or +more "dups". Each "dup" keeps reverse mapping information for a KSM +page with ``page->mapping`` pointing to that "dup". + +Every "chain" and all "dups" linked into a "chain" enforce the +invariant that they represent the same write protected memory content, +even if each "dup" will be pointed by a different KSM page copy of +that content. + +This way the stable tree lookup computational complexity is unaffected +if compared to an unlimited list of reverse mappings. It is still +enforced that there cannot be KSM page content duplicates in the +stable tree itself. + +The deduplication limit enforced by ``max_page_sharing`` is required +to avoid the virtual memory rmap lists to grow too large. The rmap +walk has O(N) complexity where N is the number of rmap_items +(i.e. virtual mappings) that are sharing the page, which is in turn +capped by ``max_page_sharing``. So this effectively spreads the linear +O(N) computational complexity from rmap walk context over different +KSM pages. The ksmd walk over the stable_node "chains" is also O(N), +but N is the number of stable_node "dups", not the number of +rmap_items, so it has not a significant impact on ksmd performance. In +practice the best stable_node "dup" candidate will be kept and found +at the head of the "dups" list. + +High values of ``max_page_sharing`` result in faster memory merging +(because there will be fewer stable_node dups queued into the +stable_node chain->hlist to check for pruning) and higher +deduplication factor at the expense of slower worst case for rmap +walks for any KSM page which can happen during swapping, compaction, +NUMA balancing and page migration. + +The ``stable_node_dups/stable_node_chains`` ratio is also affected by the +``max_page_sharing`` tunable, and an high ratio may indicate fragmentation +in the stable_node dups, which could be solved by introducing +fragmentation algorithms in ksmd which would refile rmap_items from +one stable_node dup to another stable_node dup, in order to free up +stable_node "dups" with few rmap_items in them, but that may increase +the ksmd CPU usage and possibly slowdown the readonly computations on +the KSM pages of the applications. + +The whole list of stable_node "dups" linked in the stable_node +"chains" is scanned periodically in order to prune stale stable_nodes. +The frequency of such scans is defined by +``stable_node_chains_prune_millisecs`` sysfs tunable. + +Reference +--------- +.. kernel-doc:: mm/ksm.c + :functions: mm_slot ksm_scan stable_node rmap_item + +-- +Izik Eidus, +Hugh Dickins, 17 Nov 2009 diff --git a/Documentation/vm/ksm.txt b/Documentation/vm/ksm.txt deleted file mode 100644 index 6686bd267dc9d..0000000000000 --- a/Documentation/vm/ksm.txt +++ /dev/null @@ -1,178 +0,0 @@ -How to use the Kernel Samepage Merging feature ----------------------------------------------- - -KSM is a memory-saving de-duplication feature, enabled by CONFIG_KSM=y, -added to the Linux kernel in 2.6.32. See mm/ksm.c for its implementation, -and http://lwn.net/Articles/306704/ and http://lwn.net/Articles/330589/ - -The KSM daemon ksmd periodically scans those areas of user memory which -have been registered with it, looking for pages of identical content which -can be replaced by a single write-protected page (which is automatically -copied if a process later wants to update its content). - -KSM was originally developed for use with KVM (where it was known as -Kernel Shared Memory), to fit more virtual machines into physical memory, -by sharing the data common between them. But it can be useful to any -application which generates many instances of the same data. - -KSM only merges anonymous (private) pages, never pagecache (file) pages. -KSM's merged pages were originally locked into kernel memory, but can now -be swapped out just like other user pages (but sharing is broken when they -are swapped back in: ksmd must rediscover their identity and merge again). - -KSM only operates on those areas of address space which an application -has advised to be likely candidates for merging, by using the madvise(2) -system call: int madvise(addr, length, MADV_MERGEABLE). - -The app may call int madvise(addr, length, MADV_UNMERGEABLE) to cancel -that advice and restore unshared pages: whereupon KSM unmerges whatever -it merged in that range. Note: this unmerging call may suddenly require -more memory than is available - possibly failing with EAGAIN, but more -probably arousing the Out-Of-Memory killer. - -If KSM is not configured into the running kernel, madvise MADV_MERGEABLE -and MADV_UNMERGEABLE simply fail with EINVAL. If the running kernel was -built with CONFIG_KSM=y, those calls will normally succeed: even if the -the KSM daemon is not currently running, MADV_MERGEABLE still registers -the range for whenever the KSM daemon is started; even if the range -cannot contain any pages which KSM could actually merge; even if -MADV_UNMERGEABLE is applied to a range which was never MADV_MERGEABLE. - -If a region of memory must be split into at least one new MADV_MERGEABLE -or MADV_UNMERGEABLE region, the madvise may return ENOMEM if the process -will exceed vm.max_map_count (see Documentation/sysctl/vm.txt). - -Like other madvise calls, they are intended for use on mapped areas of -the user address space: they will report ENOMEM if the specified range -includes unmapped gaps (though working on the intervening mapped areas), -and might fail with EAGAIN if not enough memory for internal structures. - -Applications should be considerate in their use of MADV_MERGEABLE, -restricting its use to areas likely to benefit. KSM's scans may use a lot -of processing power: some installations will disable KSM for that reason. - -The KSM daemon is controlled by sysfs files in /sys/kernel/mm/ksm/, -readable by all but writable only by root: - -pages_to_scan - how many present pages to scan before ksmd goes to sleep - e.g. "echo 100 > /sys/kernel/mm/ksm/pages_to_scan" - Default: 100 (chosen for demonstration purposes) - -sleep_millisecs - how many milliseconds ksmd should sleep before next scan - e.g. "echo 20 > /sys/kernel/mm/ksm/sleep_millisecs" - Default: 20 (chosen for demonstration purposes) - -merge_across_nodes - specifies if pages from different numa nodes can be merged. - When set to 0, ksm merges only pages which physically - reside in the memory area of same NUMA node. That brings - lower latency to access of shared pages. Systems with more - nodes, at significant NUMA distances, are likely to benefit - from the lower latency of setting 0. Smaller systems, which - need to minimize memory usage, are likely to benefit from - the greater sharing of setting 1 (default). You may wish to - compare how your system performs under each setting, before - deciding on which to use. merge_across_nodes setting can be - changed only when there are no ksm shared pages in system: - set run 2 to unmerge pages first, then to 1 after changing - merge_across_nodes, to remerge according to the new setting. - Default: 1 (merging across nodes as in earlier releases) - -run - set 0 to stop ksmd from running but keep merged pages, - set 1 to run ksmd e.g. "echo 1 > /sys/kernel/mm/ksm/run", - set 2 to stop ksmd and unmerge all pages currently merged, - but leave mergeable areas registered for next run - Default: 0 (must be changed to 1 to activate KSM, - except if CONFIG_SYSFS is disabled) - -use_zero_pages - specifies whether empty pages (i.e. allocated pages - that only contain zeroes) should be treated specially. - When set to 1, empty pages are merged with the kernel - zero page(s) instead of with each other as it would - happen normally. This can improve the performance on - architectures with coloured zero pages, depending on - the workload. Care should be taken when enabling this - setting, as it can potentially degrade the performance - of KSM for some workloads, for example if the checksums - of pages candidate for merging match the checksum of - an empty page. This setting can be changed at any time, - it is only effective for pages merged after the change. - Default: 0 (normal KSM behaviour as in earlier releases) - -max_page_sharing - Maximum sharing allowed for each KSM page. This - enforces a deduplication limit to avoid the virtual - memory rmap lists to grow too large. The minimum - value is 2 as a newly created KSM page will have at - least two sharers. 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 spread the 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. The higher this value the - faster KSM will merge the memory (because there - will be fewer stable_node dups queued into the - stable_node chain->hlist to check for pruning) and - the higher the deduplication factor will be, but - the slowest the worst case rmap walk could be for - any given KSM page. Slowing down the rmap_walk - means there will be higher latency for certain - virtual memory operations happening during - swapping, compaction, NUMA balancing and page - migration, in turn decreasing responsiveness for - the caller of those virtual memory operations. The - scheduler latency of other tasks not involved with - the VM operations doing the rmap walk is not - affected by this parameter as the rmap walks are - always schedule friendly themselves. - -stable_node_chains_prune_millisecs - How frequently to walk the whole - list of stable_node "dups" linked in the - stable_node "chains" in order to prune stale - stable_nodes. Smaller milllisecs values will free - up the KSM metadata with lower latency, but they - will make ksmd use more CPU during the scan. This - only applies to the stable_node chains so it's a - noop if not a single KSM page hit the - max_page_sharing yet (there would be no stable_node - chains in such case). - -The effectiveness of KSM and MADV_MERGEABLE is shown in /sys/kernel/mm/ksm/: - -pages_shared - how many shared pages are being used -pages_sharing - how many more sites are sharing them i.e. how much saved -pages_unshared - how many pages unique but repeatedly checked for merging -pages_volatile - how many pages changing too fast to be placed in a tree -full_scans - how many times all mergeable areas have been scanned - -stable_node_chains - number of stable node chains allocated, this is - effectively the number of KSM pages that hit the - max_page_sharing limit -stable_node_dups - number of stable node dups queued into the - stable_node chains - -A high ratio of pages_sharing to pages_shared indicates good sharing, but -a high ratio of pages_unshared to pages_sharing indicates wasted effort. -pages_volatile embraces several different kinds of activity, but a high -proportion there would also indicate poor use of madvise MADV_MERGEABLE. - -The maximum possible page_sharing/page_shared ratio is limited by the -max_page_sharing tunable. To increase the ratio max_page_sharing must -be increased accordingly. - -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 freeup -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. - -Izik Eidus, -Hugh Dickins, 17 Nov 2009 diff --git a/Documentation/vm/mmu_notifier.rst b/Documentation/vm/mmu_notifier.rst new file mode 100644 index 0000000000000..47baa1cf28c57 --- /dev/null +++ b/Documentation/vm/mmu_notifier.rst @@ -0,0 +1,99 @@ +.. _mmu_notifier: + +When do you need to notify inside page table lock ? +=================================================== + +When clearing a pte/pmd we are given a choice to notify the event through +(notify version of \*_clear_flush call mmu_notifier_invalidate_range) under +the page table lock. But that notification is not necessary in all cases. + +For secondary TLB (non CPU TLB) like IOMMU TLB or device TLB (when device use +thing like ATS/PASID to get the IOMMU to walk the CPU page table to access a +process virtual address space). There is only 2 cases when you need to notify +those secondary TLB while holding page table lock when clearing a pte/pmd: + + A) page backing address is free before mmu_notifier_invalidate_range_end() + B) a page table entry is updated to point to a new page (COW, write fault + on zero page, __replace_page(), ...) + +Case A is obvious you do not want to take the risk for the device to write to +a page that might now be used by some completely different task. + +Case B is more subtle. For correctness it requires the following sequence to +happen: + + - take page table lock + - clear page table entry and notify ([pmd/pte]p_huge_clear_flush_notify()) + - set page table entry to point to new page + +If clearing the page table entry is not followed by a notify before setting +the new pte/pmd value then you can break memory model like C11 or C++11 for +the device. + +Consider the following scenario (device use a feature similar to ATS/PASID): + +Two address addrA and addrB such that \|addrA - addrB\| >= PAGE_SIZE we assume +they are write protected for COW (other case of B apply too). + +:: + + [Time N] -------------------------------------------------------------------- + CPU-thread-0 {try to write to addrA} + CPU-thread-1 {try to write to addrB} + CPU-thread-2 {} + CPU-thread-3 {} + DEV-thread-0 {read addrA and populate device TLB} + DEV-thread-2 {read addrB and populate device TLB} + [Time N+1] ------------------------------------------------------------------ + CPU-thread-0 {COW_step0: {mmu_notifier_invalidate_range_start(addrA)}} + CPU-thread-1 {COW_step0: {mmu_notifier_invalidate_range_start(addrB)}} + CPU-thread-2 {} + CPU-thread-3 {} + DEV-thread-0 {} + DEV-thread-2 {} + [Time N+2] ------------------------------------------------------------------ + CPU-thread-0 {COW_step1: {update page table to point to new page for addrA}} + CPU-thread-1 {COW_step1: {update page table to point to new page for addrB}} + CPU-thread-2 {} + CPU-thread-3 {} + DEV-thread-0 {} + DEV-thread-2 {} + [Time N+3] ------------------------------------------------------------------ + CPU-thread-0 {preempted} + CPU-thread-1 {preempted} + CPU-thread-2 {write to addrA which is a write to new page} + CPU-thread-3 {} + DEV-thread-0 {} + DEV-thread-2 {} + [Time N+3] ------------------------------------------------------------------ + CPU-thread-0 {preempted} + CPU-thread-1 {preempted} + CPU-thread-2 {} + CPU-thread-3 {write to addrB which is a write to new page} + DEV-thread-0 {} + DEV-thread-2 {} + [Time N+4] ------------------------------------------------------------------ + CPU-thread-0 {preempted} + CPU-thread-1 {COW_step3: {mmu_notifier_invalidate_range_end(addrB)}} + CPU-thread-2 {} + CPU-thread-3 {} + DEV-thread-0 {} + DEV-thread-2 {} + [Time N+5] ------------------------------------------------------------------ + CPU-thread-0 {preempted} + CPU-thread-1 {} + CPU-thread-2 {} + CPU-thread-3 {} + DEV-thread-0 {read addrA from old page} + DEV-thread-2 {read addrB from new page} + +So here because at time N+2 the clear page table entry was not pair with a +notification to invalidate the secondary TLB, the device see the new value for +addrB before seing the new value for addrA. This break total memory ordering +for the device. + +When changing a pte to write protect or to point to a new write protected page +with same content (KSM) it is fine to delay the mmu_notifier_invalidate_range +call to mmu_notifier_invalidate_range_end() outside the page table lock. This +is true even if the thread doing the page table update is preempted right after +releasing page table lock but before call mmu_notifier_invalidate_range_end(). diff --git a/Documentation/vm/mmu_notifier.txt b/Documentation/vm/mmu_notifier.txt deleted file mode 100644 index 23b462566bb76..0000000000000 --- a/Documentation/vm/mmu_notifier.txt +++ /dev/null @@ -1,93 +0,0 @@ -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 seing the new value for addrA. This break total memory ordering -for the device. - -When changing a pte to write protect or to point to a new write protected page -with same content (KSM) it is fine to delay the mmu_notifier_invalidate_range -call to mmu_notifier_invalidate_range_end() outside the page table lock. This -is true even if the thread doing the page table update is preempted right after -releasing page table lock but before call mmu_notifier_invalidate_range_end(). diff --git a/Documentation/vm/numa b/Documentation/vm/numa.rst index a31b85b9bb887..185d8a5681681 100644 --- a/Documentation/vm/numa +++ b/Documentation/vm/numa.rst @@ -1,6 +1,10 @@ +.. _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. @@ -106,7 +110,7 @@ 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 Documentation/vm/numa_memory_policy.txt.] +[see Documentation/admin-guide/mm/numa_memory_policy.rst.] System administrators can restrict the CPUs and nodes' memories that a non- privileged user can specify in the scheduling or NUMA commands and functions diff --git a/Documentation/vm/numa_memory_policy.txt b/Documentation/vm/numa_memory_policy.txt deleted file mode 100644 index 622b927816e78..0000000000000 --- a/Documentation/vm/numa_memory_policy.txt +++ /dev/null @@ -1,452 +0,0 @@ - -What is Linux Memory Policy? - -In the Linux kernel, "memory policy" determines from which node the kernel will -allocate memory in a NUMA system or in an emulated NUMA system. Linux has -supported platforms with Non-Uniform Memory Access architectures since 2.4.?. -The current memory policy support was added to Linux 2.6 around May 2004. This -document attempts to describe the concepts and APIs of the 2.6 memory policy -support. - -Memory policies should not be confused with cpusets -(Documentation/cgroup-v1/cpusets.txt) -which is an administrative mechanism for restricting the nodes from which -memory may be allocated by a set of processes. Memory policies are a -programming interface that a NUMA-aware application can take advantage of. When -both cpusets and policies are applied to a task, the restrictions of the cpuset -takes priority. See "MEMORY POLICIES AND CPUSETS" below for more details. - -MEMORY POLICY CONCEPTS - -Scope of Memory Policies - -The Linux kernel supports _scopes_ of memory policy, described here from -most general to most specific: - - System Default Policy: this policy is "hard coded" into the kernel. It - is the policy that governs all page allocations that aren't controlled - by one of the more specific policy scopes discussed below. When the - system is "up and running", the system default policy will use "local - allocation" described below. However, during boot up, the system - default policy will be set to interleave allocations across all nodes - with "sufficient" memory, so as not to overload the initial boot node - with boot-time allocations. - - Task/Process Policy: this is an optional, per-task policy. When defined - for a specific task, this policy controls all page allocations made by or - on behalf of the task that aren't controlled by a more specific scope. - If a task does not define a task policy, then all page allocations that - would have been controlled by the task policy "fall back" to the System - Default Policy. - - The task policy applies to the entire address space of a task. Thus, - it is inheritable, and indeed is inherited, across both fork() - [clone() w/o the CLONE_VM flag] and exec*(). This allows a parent task - to establish the task policy for a child task exec()'d from an - executable image that has no awareness of memory policy. See the - MEMORY POLICY APIS section, below, for an overview of the system call - that a task may use to set/change its task/process policy. - - In a multi-threaded task, task policies apply only to the thread - [Linux kernel task] that installs the policy and any threads - subsequently created by that thread. Any sibling threads existing - at the time a new task policy is installed retain their current - policy. - - A task policy applies only to pages allocated after the policy is - installed. Any pages already faulted in by the task when the task - changes its task policy remain where they were allocated based on - the policy at the time they were allocated. - - VMA Policy: A "VMA" or "Virtual Memory Area" refers to a range of a task's - virtual address space. A task may define a specific policy for a range - of its virtual address space. See the MEMORY POLICIES APIS section, - below, for an overview of the mbind() system call used to set a VMA - policy. - - A VMA policy will govern the allocation of pages that back this region of - the address space. Any regions of the task's address space that don't - have an explicit VMA policy will fall back to the task policy, which may - itself fall back to the System Default Policy. - - VMA policies have a few complicating details: - - VMA policy applies ONLY to anonymous pages. These include pages - allocated for anonymous segments, such as the task stack and heap, and - any regions of the address space mmap()ed with the MAP_ANONYMOUS flag. - If a VMA policy is applied to a file mapping, it will be ignored if - the mapping used the MAP_SHARED flag. If the file mapping used the - MAP_PRIVATE flag, the VMA policy will only be applied when an - anonymous page is allocated on an attempt to write to the mapping-- - i.e., at Copy-On-Write. - - VMA policies are shared between all tasks that share a virtual address - space--a.k.a. threads--independent of when the policy is installed; and - they are inherited across fork(). However, because VMA policies refer - to a specific region of a task's address space, and because the address - space is discarded and recreated on exec*(), VMA policies are NOT - inheritable across exec(). Thus, only NUMA-aware applications may - use VMA policies. - - A task may install a new VMA policy on a sub-range of a previously - mmap()ed region. When this happens, Linux splits the existing virtual - memory area into 2 or 3 VMAs, each with it's own policy. - - By default, VMA policy applies only to pages allocated after the policy - is installed. Any pages already faulted into the VMA range remain - where they were allocated based on the policy at the time they were - allocated. However, since 2.6.16, Linux supports page migration via - the mbind() system call, so that page contents can be moved to match - a newly installed policy. - - Shared Policy: Conceptually, shared policies apply to "memory objects" - mapped shared into one or more tasks' distinct address spaces. An - application installs a shared policies the same way as VMA policies--using - the mbind() system call specifying a range of virtual addresses that map - the shared object. However, unlike VMA policies, which can be considered - to be an attribute of a range of a task's address space, shared policies - apply directly to the shared object. Thus, all tasks that attach to the - object share the policy, and all pages allocated for the shared object, - by any task, will obey the shared policy. - - As of 2.6.22, only shared memory segments, created by shmget() or - mmap(MAP_ANONYMOUS|MAP_SHARED), support shared policy. When shared - policy support was added to Linux, the associated data structures were - added to hugetlbfs shmem segments. At the time, hugetlbfs did not - support allocation at fault time--a.k.a lazy allocation--so hugetlbfs - shmem segments were never "hooked up" to the shared policy support. - Although hugetlbfs segments now support lazy allocation, their support - for shared policy has not been completed. - - As mentioned above [re: VMA policies], allocations of page cache - pages for regular files mmap()ed with MAP_SHARED ignore any VMA - policy installed on the virtual address range backed by the shared - file mapping. Rather, shared page cache pages, including pages backing - private mappings that have not yet been written by the task, follow - task policy, if any, else System Default Policy. - - The shared policy infrastructure supports different policies on subset - ranges of the shared object. However, Linux still splits the VMA of - the task that installs the policy for each range of distinct policy. - Thus, different tasks that attach to a shared memory segment can have - different VMA configurations mapping that one shared object. This - can be seen by examining the /proc/<pid>/numa_maps of tasks sharing - a shared memory region, when one task has installed shared policy on - one or more ranges of the region. - -Components of Memory Policies - - A Linux memory policy consists of a "mode", optional mode flags, and an - optional set of nodes. The mode determines the behavior of the policy, - the optional mode flags determine the behavior of the mode, and the - optional set of nodes can be viewed as the arguments to the policy - behavior. - - Internally, memory policies are implemented by a reference counted - structure, struct mempolicy. Details of this structure will be discussed - in context, below, as required to explain the behavior. - - Linux memory policy supports the following 4 behavioral modes: - - Default Mode--MPOL_DEFAULT: This mode is only used in the memory - policy APIs. Internally, MPOL_DEFAULT is converted to the NULL - memory policy in all policy scopes. Any existing non-default policy - will simply be removed when MPOL_DEFAULT is specified. As a result, - MPOL_DEFAULT means "fall back to the next most specific policy scope." - - For example, a NULL or default task policy will fall back to the - system default policy. A NULL or default vma policy will fall - back to the task policy. - - When specified in one of the memory policy APIs, the Default mode - does not use the optional set of nodes. - - It is an error for the set of nodes specified for this policy to - be non-empty. - - MPOL_BIND: This mode specifies that memory must come from the - set of nodes specified by the policy. Memory will be allocated from - the node in the set with sufficient free memory that is closest to - the node where the allocation takes place. - - MPOL_PREFERRED: This mode specifies that the allocation should be - attempted from the single node specified in the policy. If that - allocation fails, the kernel will search other nodes, in order of - increasing distance from the preferred node based on information - provided by the platform firmware. - - Internally, the Preferred policy uses a single node--the - preferred_node member of struct mempolicy. When the internal - mode flag MPOL_F_LOCAL is set, the preferred_node is ignored and - the policy is interpreted as local allocation. "Local" allocation - policy can be viewed as a Preferred policy that starts at the node - containing the cpu where the allocation takes place. - - It is possible for the user to specify that local allocation is - always preferred by passing an empty nodemask with this mode. - If an empty nodemask is passed, the policy cannot use the - MPOL_F_STATIC_NODES or MPOL_F_RELATIVE_NODES flags described - below. - - MPOL_INTERLEAVED: This mode specifies that page allocations be - interleaved, on a page granularity, across the nodes specified in - the policy. This mode also behaves slightly differently, based on - the context where it is used: - - For allocation of anonymous pages and shared memory pages, - Interleave mode indexes the set of nodes specified by the policy - using the page offset of the faulting address into the segment - [VMA] containing the address modulo the number of nodes specified - by the policy. It then attempts to allocate a page, starting at - the selected node, as if the node had been specified by a Preferred - policy or had been selected by a local allocation. That is, - allocation will follow the per node zonelist. - - For allocation of page cache pages, Interleave mode indexes the set - of nodes specified by the policy using a node counter maintained - per task. This counter wraps around to the lowest specified node - after it reaches the highest specified node. This will tend to - spread the pages out over the nodes specified by the policy based - on the order in which they are allocated, rather than based on any - page offset into an address range or file. During system boot up, - the temporary interleaved system default policy works in this - mode. - - Linux memory policy supports the following optional mode flags: - - MPOL_F_STATIC_NODES: This flag specifies that the nodemask passed by - the user should not be remapped if the task or VMA's set of allowed - nodes changes after the memory policy has been defined. - - Without this flag, anytime a mempolicy is rebound because of a - change in the set of allowed nodes, the node (Preferred) or - nodemask (Bind, Interleave) is remapped to the new set of - allowed nodes. This may result in nodes being used that were - previously undesired. - - With this flag, if the user-specified nodes overlap with the - nodes allowed by the task's cpuset, then the memory policy is - applied to their intersection. If the two sets of nodes do not - overlap, the Default policy is used. - - For example, consider a task that is attached to a cpuset with - mems 1-3 that sets an Interleave policy over the same set. If - the cpuset's mems change to 3-5, the Interleave will now occur - over nodes 3, 4, and 5. With this flag, however, since only node - 3 is allowed from the user's nodemask, the "interleave" only - occurs over that node. If no nodes from the user's nodemask are - now allowed, the Default behavior is used. - - MPOL_F_STATIC_NODES cannot be combined with the - MPOL_F_RELATIVE_NODES flag. It also cannot be used for - MPOL_PREFERRED policies that were created with an empty nodemask - (local allocation). - - MPOL_F_RELATIVE_NODES: This flag specifies that the nodemask passed - by the user will be mapped relative to the set of the task or VMA's - set of allowed nodes. The kernel stores the user-passed nodemask, - and if the allowed nodes changes, then that original nodemask will - be remapped relative to the new set of allowed nodes. - - Without this flag (and without MPOL_F_STATIC_NODES), anytime a - mempolicy is rebound because of a change in the set of allowed - nodes, the node (Preferred) or nodemask (Bind, Interleave) is - remapped to the new set of allowed nodes. That remap may not - preserve the relative nature of the user's passed nodemask to its - set of allowed nodes upon successive rebinds: a nodemask of - 1,3,5 may be remapped to 7-9 and then to 1-3 if the set of - allowed nodes is restored to its original state. - - With this flag, the remap is done so that the node numbers from - the user's passed nodemask are relative to the set of allowed - nodes. In other words, if nodes 0, 2, and 4 are set in the user's - nodemask, the policy will be effected over the first (and in the - Bind or Interleave case, the third and fifth) nodes in the set of - allowed nodes. The nodemask passed by the user represents nodes - relative to task or VMA's set of allowed nodes. - - If the user's nodemask includes nodes that are outside the range - of the new set of allowed nodes (for example, node 5 is set in - the user's nodemask when the set of allowed nodes is only 0-3), - then the remap wraps around to the beginning of the nodemask and, - if not already set, sets the node in the mempolicy nodemask. - - For example, consider a task that is attached to a cpuset with - mems 2-5 that sets an Interleave policy over the same set with - MPOL_F_RELATIVE_NODES. If the cpuset's mems change to 3-7, the - interleave now occurs over nodes 3,5-7. If the cpuset's mems - then change to 0,2-3,5, then the interleave occurs over nodes - 0,2-3,5. - - Thanks to the consistent remapping, applications preparing - nodemasks to specify memory policies using this flag should - disregard their current, actual cpuset imposed memory placement - and prepare the nodemask as if they were always located on - memory nodes 0 to N-1, where N is the number of memory nodes the - policy is intended to manage. Let the kernel then remap to the - set of memory nodes allowed by the task's cpuset, as that may - change over time. - - MPOL_F_RELATIVE_NODES cannot be combined with the - MPOL_F_STATIC_NODES flag. It also cannot be used for - MPOL_PREFERRED policies that were created with an empty nodemask - (local allocation). - -MEMORY POLICY REFERENCE COUNTING - -To resolve use/free races, struct mempolicy contains an atomic reference -count field. Internal interfaces, mpol_get()/mpol_put() increment and -decrement this reference count, respectively. mpol_put() will only free -the structure back to the mempolicy kmem cache when the reference count -goes to zero. - -When a new memory policy is allocated, its reference count is initialized -to '1', representing the reference held by the task that is installing the -new policy. When a pointer to a memory policy structure is stored in another -structure, another reference is added, as the task's reference will be dropped -on completion of the policy installation. - -During run-time "usage" of the policy, we attempt to minimize atomic operations -on the reference count, as this can lead to cache lines bouncing between cpus -and NUMA nodes. "Usage" here means one of the following: - -1) querying of the policy, either by the task itself [using the get_mempolicy() - API discussed below] or by another task using the /proc/<pid>/numa_maps - interface. - -2) examination of the policy to determine the policy mode and associated node - or node lists, if any, for page allocation. This is considered a "hot - path". Note that for MPOL_BIND, the "usage" extends across the entire - allocation process, which may sleep during page reclaimation, because the - BIND policy nodemask is used, by reference, to filter ineligible nodes. - -We can avoid taking an extra reference during the usages listed above as -follows: - -1) we never need to get/free the system default policy as this is never - changed nor freed, once the system is up and running. - -2) for querying the policy, we do not need to take an extra reference on the - target task's task policy nor vma policies because we always acquire the - task's mm's mmap_sem for read during the query. The set_mempolicy() and - mbind() APIs [see below] always acquire the mmap_sem for write when - installing or replacing task or vma policies. Thus, there is no possibility - of a task or thread freeing a policy while another task or thread is - querying it. - -3) Page allocation usage of task or vma policy occurs in the fault path where - we hold them mmap_sem for read. Again, because replacing the task or vma - policy requires that the mmap_sem be held for write, the policy can't be - freed out from under us while we're using it for page allocation. - -4) Shared policies require special consideration. One task can replace a - shared memory policy while another task, with a distinct mmap_sem, is - querying or allocating a page based on the policy. To resolve this - potential race, the shared policy infrastructure adds an extra reference - to the shared policy during lookup while holding a spin lock on the shared - policy management structure. This requires that we drop this extra - reference when we're finished "using" the policy. We must drop the - extra reference on shared policies in the same query/allocation paths - used for non-shared policies. For this reason, shared policies are marked - as such, and the extra reference is dropped "conditionally"--i.e., only - for shared policies. - - Because of this extra reference counting, and because we must lookup - shared policies in a tree structure under spinlock, shared policies are - more expensive to use in the page allocation path. This is especially - true for shared policies on shared memory regions shared by tasks running - on different NUMA nodes. This extra overhead can be avoided by always - falling back to task or system default policy for shared memory regions, - or by prefaulting the entire shared memory region into memory and locking - it down. However, this might not be appropriate for all applications. - -MEMORY POLICY APIs - -Linux supports 3 system calls for controlling memory policy. These APIS -always affect only the calling task, the calling task's address space, or -some shared object mapped into the calling task's address space. - - Note: the headers that define these APIs and the parameter data types - for user space applications reside in a package that is not part of - the Linux kernel. The kernel system call interfaces, with the 'sys_' - prefix, are defined in <linux/syscalls.h>; the mode and flag - definitions are defined in <linux/mempolicy.h>. - -Set [Task] Memory Policy: - - long set_mempolicy(int mode, const unsigned long *nmask, - unsigned long maxnode); - - Set's the calling task's "task/process memory policy" to mode - specified by the 'mode' argument and the set of nodes defined - by 'nmask'. 'nmask' points to a bit mask of node ids containing - at least 'maxnode' ids. Optional mode flags may be passed by - combining the 'mode' argument with the flag (for example: - MPOL_INTERLEAVE | MPOL_F_STATIC_NODES). - - See the set_mempolicy(2) man page for more details - - -Get [Task] Memory Policy or Related Information - - long get_mempolicy(int *mode, - const unsigned long *nmask, unsigned long maxnode, - void *addr, int flags); - - Queries the "task/process memory policy" of the calling task, or - the policy or location of a specified virtual address, depending - on the 'flags' argument. - - See the get_mempolicy(2) man page for more details - - -Install VMA/Shared Policy for a Range of Task's Address Space - - long mbind(void *start, unsigned long len, int mode, - const unsigned long *nmask, unsigned long maxnode, - unsigned flags); - - mbind() installs the policy specified by (mode, nmask, maxnodes) as - a VMA policy for the range of the calling task's address space - specified by the 'start' and 'len' arguments. Additional actions - may be requested via the 'flags' argument. - - See the mbind(2) man page for more details. - -MEMORY POLICY COMMAND LINE INTERFACE - -Although not strictly part of the Linux implementation of memory policy, -a command line tool, numactl(8), exists that allows one to: - -+ set the task policy for a specified program via set_mempolicy(2), fork(2) and - exec(2) - -+ set the shared policy for a shared memory segment via mbind(2) - -The numactl(8) tool is packaged with the run-time version of the library -containing the memory policy system call wrappers. Some distributions -package the headers and compile-time libraries in a separate development -package. - - -MEMORY POLICIES AND CPUSETS - -Memory policies work within cpusets as described above. For memory policies -that require a node or set of nodes, the nodes are restricted to the set of -nodes whose memories are allowed by the cpuset constraints. If the nodemask -specified for the policy contains nodes that are not allowed by the cpuset and -MPOL_F_RELATIVE_NODES is not used, the intersection of the set of nodes -specified for the policy and the set of nodes with memory is used. If the -result is the empty set, the policy is considered invalid and cannot be -installed. If MPOL_F_RELATIVE_NODES is used, the policy's nodes are mapped -onto and folded into the task's set of allowed nodes as previously described. - -The interaction of memory policies and cpusets can be problematic when tasks -in two cpusets share access to a memory region, such as shared memory segments -created by shmget() of mmap() with the MAP_ANONYMOUS and MAP_SHARED flags, and -any of the tasks install shared policy on the region, only nodes whose -memories are allowed in both cpusets may be used in the policies. Obtaining -this information requires "stepping outside" the memory policy APIs to use the -cpuset information and requires that one know in what cpusets other task might -be attaching to the shared region. Furthermore, if the cpusets' allowed -memory sets are disjoint, "local" allocation is the only valid policy. diff --git a/Documentation/vm/overcommit-accounting b/Documentation/vm/overcommit-accounting deleted file mode 100644 index cbfaaa674118d..0000000000000 --- a/Documentation/vm/overcommit-accounting +++ /dev/null @@ -1,80 +0,0 @@ -The Linux kernel supports the following overcommit handling modes - -0 - Heuristic overcommit handling. Obvious overcommits of - address space are refused. Used for a typical system. It - ensures a seriously wild allocation fails while allowing - overcommit to reduce swap usage. root is allowed to - allocate slightly more memory in this mode. This is the - default. - -1 - Always overcommit. Appropriate for some scientific - applications. Classic example is code using sparse arrays - and just relying on the virtual memory consisting almost - entirely of zero pages. - -2 - Don't overcommit. The total address space commit - for the system is not permitted to exceed swap + a - configurable amount (default is 50%) of physical RAM. - Depending on the amount you use, in most situations - this means a process will not be killed while accessing - pages but will receive errors on memory allocation as - appropriate. - - Useful for applications that want to guarantee their - memory allocations will be available in the future - without having to initialize every page. - -The overcommit policy is set via the sysctl `vm.overcommit_memory'. - -The overcommit amount can be set via `vm.overcommit_ratio' (percentage) -or `vm.overcommit_kbytes' (absolute value). - -The current overcommit limit and amount committed are viewable in -/proc/meminfo as CommitLimit and Committed_AS respectively. - -Gotchas -------- - -The C language stack growth does an implicit mremap. If you want absolute -guarantees and run close to the edge you MUST mmap your stack for the -largest size you think you will need. For typical stack usage this does -not matter much but it's a corner case if you really really care - -In mode 2 the MAP_NORESERVE flag is ignored. - - -How It Works ------------- - -The overcommit is based on the following rules - -For a file backed map - SHARED or READ-only - 0 cost (the file is the map not swap) - PRIVATE WRITABLE - size of mapping per instance - -For an anonymous or /dev/zero map - SHARED - size of mapping - PRIVATE READ-only - 0 cost (but of little use) - PRIVATE WRITABLE - size of mapping per instance - -Additional accounting - Pages made writable copies by mmap - shmfs memory drawn from the same pool - -Status ------- - -o We account mmap memory mappings -o We account mprotect changes in commit -o We account mremap changes in size -o We account brk -o We account munmap -o We report the commit status in /proc -o Account and check on fork -o Review stack handling/building on exec -o SHMfs accounting -o Implement actual limit enforcement - -To Do ------ -o Account ptrace pages (this is hard) diff --git a/Documentation/vm/overcommit-accounting.rst b/Documentation/vm/overcommit-accounting.rst new file mode 100644 index 0000000000000..0dd54bbe4afa5 --- /dev/null +++ b/Documentation/vm/overcommit-accounting.rst @@ -0,0 +1,87 @@ +.. _overcommit_accounting: + +===================== +Overcommit Accounting +===================== + +The Linux kernel supports the following overcommit handling modes + +0 + Heuristic overcommit handling. Obvious overcommits of address + space are refused. Used for a typical system. It ensures a + seriously wild allocation fails while allowing overcommit to + reduce swap usage. root is allowed to allocate slightly more + memory in this mode. This is the default. + +1 + Always overcommit. Appropriate for some scientific + applications. Classic example is code using sparse arrays and + just relying on the virtual memory consisting almost entirely + of zero pages. + +2 + Don't overcommit. The total address space commit for the + system is not permitted to exceed swap + a configurable amount + (default is 50%) of physical RAM. Depending on the amount you + use, in most situations this means a process will not be + killed while accessing pages but will receive errors on memory + allocation as appropriate. + + Useful for applications that want to guarantee their memory + allocations will be available in the future without having to + initialize every page. + +The overcommit policy is set via the sysctl ``vm.overcommit_memory``. + +The overcommit amount can be set via ``vm.overcommit_ratio`` (percentage) +or ``vm.overcommit_kbytes`` (absolute value). + +The current overcommit limit and amount committed are viewable in +``/proc/meminfo`` as CommitLimit and Committed_AS respectively. + +Gotchas +======= + +The C language stack growth does an implicit mremap. If you want absolute +guarantees and run close to the edge you MUST mmap your stack for the +largest size you think you will need. For typical stack usage this does +not matter much but it's a corner case if you really really care + +In mode 2 the MAP_NORESERVE flag is ignored. + + +How It Works +============ + +The overcommit is based on the following rules + +For a file backed map + | SHARED or READ-only - 0 cost (the file is the map not swap) + | PRIVATE WRITABLE - size of mapping per instance + +For an anonymous or ``/dev/zero`` map + | SHARED - size of mapping + | PRIVATE READ-only - 0 cost (but of little use) + | PRIVATE WRITABLE - size of mapping per instance + +Additional accounting + | Pages made writable copies by mmap + | shmfs memory drawn from the same pool + +Status +====== + +* We account mmap memory mappings +* We account mprotect changes in commit +* We account mremap changes in size +* We account brk +* We account munmap +* We report the commit status in /proc +* Account and check on fork +* Review stack handling/building on exec +* SHMfs accounting +* Implement actual limit enforcement + +To Do +===== +* Account ptrace pages (this is hard) diff --git a/Documentation/vm/page_frags b/Documentation/vm/page_frags.rst index a6714565dbf9e..637cc49d1b2ff 100644 --- a/Documentation/vm/page_frags +++ b/Documentation/vm/page_frags.rst @@ -1,5 +1,8 @@ +.. _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 diff --git a/Documentation/vm/page_migration b/Documentation/vm/page_migration.rst index 496868072e24e..f68d61335abb6 100644 --- a/Documentation/vm/page_migration +++ b/Documentation/vm/page_migration.rst @@ -1,5 +1,8 @@ +.. _page_migration: + +============== Page migration --------------- +============== Page migration allows the moving of the physical location of pages between nodes in a numa system while the process is running. This means that the @@ -20,7 +23,7 @@ Page migration functions are provided by the numactl package by Andi Kleen (a version later than 0.9.3 is required. Get it from ftp://oss.sgi.com/www/projects/libnuma/download/). 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 +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. @@ -56,8 +59,8 @@ 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. -A. In kernel use of migrate_pages() ------------------------------------ +In kernel use of migrate_pages() +================================ 1. Remove pages from the LRU. @@ -78,8 +81,8 @@ A. In kernel use of migrate_pages() the new page for each page that is considered for moving. -B. How migrate_pages() works ----------------------------- +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 @@ -142,8 +145,8 @@ Steps: 20. The new page is moved to the LRU and can be scanned by the swapper etc again. -C. Non-LRU page migration -------------------------- +Non-LRU page migration +====================== Although original migration aimed for reducing the latency of memory access for NUMA, compaction who want to create high-order page is also main customer. @@ -164,89 +167,91 @@ migration path. If a driver want to make own pages movable, it should define three functions which are function pointers of struct address_space_operations. -1. bool (*isolate_page) (struct page *page, isolate_mode_t mode); +1. ``bool (*isolate_page) (struct page *page, isolate_mode_t mode);`` -What VM expects on isolate_page function of driver is to return *true* -if driver isolates page successfully. On returing true, VM marks the page -as PG_isolated so concurrent isolation in several CPUs skip the page -for isolation. If a driver cannot isolate the page, it should return *false*. + What VM expects on isolate_page function of driver is to return *true* + if driver isolates page successfully. On returing true, VM marks the page + as PG_isolated so concurrent isolation in several CPUs skip the page + for isolation. If a driver cannot isolate the page, it should return *false*. -Once page is successfully isolated, VM uses page.lru fields so driver -shouldn't expect to preserve values in that fields. + Once page is successfully isolated, VM uses page.lru fields so driver + shouldn't expect to preserve values in that fields. -2. int (*migratepage) (struct address_space *mapping, - struct page *newpage, struct page *oldpage, enum migrate_mode); +2. ``int (*migratepage) (struct address_space *mapping,`` +| ``struct page *newpage, struct page *oldpage, enum migrate_mode);`` -After isolation, VM calls migratepage of driver with isolated page. -The function of migratepage is to move content of the old page to new page -and set up fields of struct page newpage. Keep in mind that you should -indicate to the VM the oldpage is no longer movable via __ClearPageMovable() -under page_lock if you migrated the oldpage successfully and returns -MIGRATEPAGE_SUCCESS. If driver cannot migrate the page at the moment, driver -can return -EAGAIN. On -EAGAIN, VM will retry page migration in a short time -because VM interprets -EAGAIN as "temporal migration failure". On returning -any error except -EAGAIN, VM will give up the page migration without retrying -in this time. + After isolation, VM calls migratepage of driver with isolated page. + The function of migratepage is to move content of the old page to new page + and set up fields of struct page newpage. Keep in mind that you should + indicate to the VM the oldpage is no longer movable via __ClearPageMovable() + under page_lock if you migrated the oldpage successfully and returns + MIGRATEPAGE_SUCCESS. If driver cannot migrate the page at the moment, driver + can return -EAGAIN. On -EAGAIN, VM will retry page migration in a short time + because VM interprets -EAGAIN as "temporal migration failure". On returning + any error except -EAGAIN, VM will give up the page migration without retrying + in this time. -Driver shouldn't touch page.lru field VM using in the functions. + Driver shouldn't touch page.lru field VM using in the functions. -3. void (*putback_page)(struct page *); +3. ``void (*putback_page)(struct page *);`` -If migration fails on isolated page, VM should return the isolated page -to the driver so VM calls driver's putback_page with migration failed page. -In this function, driver should put the isolated page back to the own data -structure. + If migration fails on isolated page, VM should return the isolated page + to the driver so VM calls driver's putback_page with migration failed page. + In this function, driver should put the isolated page back to the own data + structure. 4. non-lru movable page flags -There are two page flags for supporting non-lru movable page. + There are two page flags for supporting non-lru movable page. -* PG_movable + * PG_movable -Driver should use the below function to make page movable under page_lock. + Driver should use the below function to make page movable under page_lock:: void __SetPageMovable(struct page *page, struct address_space *mapping) -It needs argument of address_space for registering migration family functions -which will be called by VM. Exactly speaking, PG_movable is not a real flag of -struct page. Rather than, VM reuses page->mapping's lower bits to represent it. + It needs argument of address_space for registering migration + family functions which will be called by VM. Exactly speaking, + PG_movable is not a real flag of struct page. Rather than, VM + reuses page->mapping's lower bits to represent it. +:: #define PAGE_MAPPING_MOVABLE 0x2 page->mapping = page->mapping | PAGE_MAPPING_MOVABLE; -so driver shouldn't access page->mapping directly. Instead, driver should -use page_mapping which mask off the low two bits of page->mapping under -page lock so it can get right struct address_space. - -For testing of non-lru movable page, VM supports __PageMovable function. -However, it doesn't guarantee to identify non-lru movable page because -page->mapping field is unified with other variables in struct page. -As well, if driver releases the page after isolation by VM, page->mapping -doesn't have stable value although it has PAGE_MAPPING_MOVABLE -(Look at __ClearPageMovable). But __PageMovable is cheap to catch whether -page is LRU or non-lru movable once the page has been isolated. Because -LRU pages never can have PAGE_MAPPING_MOVABLE in page->mapping. It is also -good for just peeking to test non-lru movable pages before more expensive -checking with lock_page in pfn scanning to select victim. - -For guaranteeing non-lru movable page, VM provides PageMovable function. -Unlike __PageMovable, PageMovable functions validates page->mapping and -mapping->a_ops->isolate_page under lock_page. The lock_page prevents sudden -destroying of page->mapping. - -Driver using __SetPageMovable should clear the flag via __ClearMovablePage -under page_lock before the releasing the page. - -* PG_isolated - -To prevent concurrent isolation among several CPUs, VM marks isolated page -as PG_isolated under lock_page. So if a CPU encounters PG_isolated non-lru -movable page, it can skip it. Driver doesn't need to manipulate the flag -because VM will set/clear it automatically. Keep in mind that if driver -sees PG_isolated page, it means the page have been isolated by VM so it -shouldn't touch page.lru field. -PG_isolated is alias with PG_reclaim flag so driver shouldn't use the flag -for own purpose. + so driver shouldn't access page->mapping directly. Instead, driver should + use page_mapping which mask off the low two bits of page->mapping under + page lock so it can get right struct address_space. + + For testing of non-lru movable page, VM supports __PageMovable function. + However, it doesn't guarantee to identify non-lru movable page because + page->mapping field is unified with other variables in struct page. + As well, if driver releases the page after isolation by VM, page->mapping + doesn't have stable value although it has PAGE_MAPPING_MOVABLE + (Look at __ClearPageMovable). But __PageMovable is cheap to catch whether + page is LRU or non-lru movable once the page has been isolated. Because + LRU pages never can have PAGE_MAPPING_MOVABLE in page->mapping. It is also + good for just peeking to test non-lru movable pages before more expensive + checking with lock_page in pfn scanning to select victim. + + For guaranteeing non-lru movable page, VM provides PageMovable function. + Unlike __PageMovable, PageMovable functions validates page->mapping and + mapping->a_ops->isolate_page under lock_page. The lock_page prevents sudden + destroying of page->mapping. + + Driver using __SetPageMovable should clear the flag via __ClearMovablePage + under page_lock before the releasing the page. + + * PG_isolated + + To prevent concurrent isolation among several CPUs, VM marks isolated page + as PG_isolated under lock_page. So if a CPU encounters PG_isolated non-lru + movable page, it can skip it. Driver doesn't need to manipulate the flag + because VM will set/clear it automatically. Keep in mind that if driver + sees PG_isolated page, it means the page have been isolated by VM so it + shouldn't touch page.lru field. + PG_isolated is alias with PG_reclaim flag so driver shouldn't use the flag + for own purpose. Christoph Lameter, May 8, 2006. Minchan Kim, Mar 28, 2016. diff --git a/Documentation/vm/page_owner.txt b/Documentation/vm/page_owner.rst index ffff1439076a5..0ed5ab8c7ab46 100644 --- a/Documentation/vm/page_owner.txt +++ b/Documentation/vm/page_owner.rst @@ -1,7 +1,11 @@ +.. _page_owner: + +================================================== page owner: Tracking about who allocated each page ------------------------------------------------------------ +================================================== -* Introduction +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. @@ -34,13 +38,15 @@ not affect to allocation performance, especially if the static keys jump label patching functionality is available. Following is the kernel's code size change due to this facility. -- Without page owner +- Without page owner:: + text data bss dec hex filename - 40662 1493 644 42799 a72f mm/page_alloc.o + 40662 1493 644 42799 a72f mm/page_alloc.o + +- With page owner:: -- With page owner text data bss dec hex filename - 40892 1493 644 43029 a815 mm/page_alloc.o + 40892 1493 644 43029 a815 mm/page_alloc.o 1427 24 8 1459 5b3 mm/page_ext.o 2722 50 0 2772 ad4 mm/page_owner.o @@ -62,21 +68,23 @@ 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 +Usage +===== + +1) Build user-space helper:: -1) Build user-space helper cd tools/vm make page_owner_sort -2) Enable page owner - Add "page_owner=on" to boot cmdline. +2) Enable page owner: add "page_owner=on" to boot cmdline. 3) Do the job what you want to debug -4) Analyze information from page owner +4) Analyze information from page owner:: + cat /sys/kernel/debug/page_owner > page_owner_full.txt grep -v ^PFN page_owner_full.txt > page_owner.txt ./page_owner_sort page_owner.txt sorted_page_owner.txt - See the result about who allocated each page - in the sorted_page_owner.txt. + See the result about who allocated each page + in the ``sorted_page_owner.txt``. diff --git a/Documentation/vm/pagemap.txt b/Documentation/vm/pagemap.txt deleted file mode 100644 index eafcefa152611..0000000000000 --- a/Documentation/vm/pagemap.txt +++ /dev/null @@ -1,183 +0,0 @@ -pagemap, from the userspace perspective ---------------------------------------- - -pagemap is a new (as of 2.6.25) set of interfaces in the kernel that allow -userspace programs to examine the page tables and related information by -reading files in /proc. - -There are four components to pagemap: - - * /proc/pid/pagemap. This file lets a userspace process find out which - physical frame each virtual page is mapped to. It contains one 64-bit - value for each virtual page, containing the following data (from - fs/proc/task_mmu.c, above pagemap_read): - - * Bits 0-54 page frame number (PFN) if present - * Bits 0-4 swap type if swapped - * Bits 5-54 swap offset if swapped - * Bit 55 pte is soft-dirty (see Documentation/vm/soft-dirty.txt) - * Bit 56 page exclusively mapped (since 4.2) - * Bits 57-60 zero - * Bit 61 page is file-page or shared-anon (since 3.5) - * Bit 62 page swapped - * Bit 63 page present - - Since Linux 4.0 only users with the CAP_SYS_ADMIN capability can get PFNs. - In 4.0 and 4.1 opens by unprivileged fail with -EPERM. Starting from - 4.2 the PFN field is zeroed if the user does not have CAP_SYS_ADMIN. - Reason: information about PFNs helps in exploiting Rowhammer vulnerability. - - If the page is not present but in swap, then the PFN contains an - encoding of the swap file number and the page's offset into the - swap. Unmapped pages return a null PFN. This allows determining - precisely which pages are mapped (or in swap) and comparing mapped - pages between processes. - - Efficient users of this interface will use /proc/pid/maps to - determine which areas of memory are actually mapped and llseek to - skip over unmapped regions. - - * /proc/kpagecount. This file contains a 64-bit count of the number of - times each page is mapped, indexed by PFN. - - * /proc/kpageflags. This file contains a 64-bit set of flags for each - page, indexed by PFN. - - The flags are (from fs/proc/page.c, above kpageflags_read): - - 0. LOCKED - 1. ERROR - 2. REFERENCED - 3. UPTODATE - 4. DIRTY - 5. LRU - 6. ACTIVE - 7. SLAB - 8. WRITEBACK - 9. RECLAIM - 10. BUDDY - 11. MMAP - 12. ANON - 13. SWAPCACHE - 14. SWAPBACKED - 15. COMPOUND_HEAD - 16. COMPOUND_TAIL - 17. HUGE - 18. UNEVICTABLE - 19. HWPOISON - 20. NOPAGE - 21. KSM - 22. THP - 23. BALLOON - 24. ZERO_PAGE - 25. IDLE - - * /proc/kpagecgroup. This file contains a 64-bit inode number of the - memory cgroup each page is charged to, indexed by PFN. Only available when - CONFIG_MEMCG is set. - -Short descriptions to the page flags: - - 0. LOCKED - page is being locked for exclusive access, eg. by undergoing read/write IO - - 7. SLAB - page is managed by the SLAB/SLOB/SLUB/SLQB kernel memory allocator - When compound page is used, SLUB/SLQB will only set this flag on the head - page; SLOB will not flag it at all. - -10. BUDDY - a free memory block managed by the buddy system allocator - The buddy system organizes free memory in blocks of various orders. - An order N block has 2^N physically contiguous pages, with the BUDDY flag - set for and _only_ for the first page. - -15. COMPOUND_HEAD -16. COMPOUND_TAIL - A compound page with order N consists of 2^N physically contiguous pages. - A compound page with order 2 takes the form of "HTTT", where H donates its - head page and T donates its tail page(s). The major consumers of compound - pages are hugeTLB pages (Documentation/vm/hugetlbpage.txt), the SLUB etc. - memory allocators and various device drivers. However in this interface, - only huge/giga pages are made visible to end users. -17. HUGE - this is an integral part of a HugeTLB page - -19. HWPOISON - hardware detected memory corruption on this page: don't touch the data! - -20. NOPAGE - no page frame exists at the requested address - -21. KSM - identical memory pages dynamically shared between one or more processes - -22. THP - contiguous pages which construct transparent hugepages - -23. BALLOON - balloon compaction page - -24. ZERO_PAGE - zero page for pfn_zero or huge_zero page - -25. IDLE - page has not been accessed since it was marked idle (see - Documentation/vm/idle_page_tracking.txt). Note that this flag may be - stale in case the page was accessed via a PTE. To make sure the flag - is up-to-date one has to read /sys/kernel/mm/page_idle/bitmap first. - - [IO related page flags] - 1. ERROR IO error occurred - 3. UPTODATE page has up-to-date data - ie. for file backed page: (in-memory data revision >= on-disk one) - 4. DIRTY page has been written to, hence contains new data - ie. for file backed page: (in-memory data revision > on-disk one) - 8. WRITEBACK page is being synced to disk - - [LRU related page flags] - 5. LRU page is in one of the LRU lists - 6. ACTIVE page is in the active LRU list -18. UNEVICTABLE page is in the unevictable (non-)LRU list - It is somehow pinned and not a candidate for LRU page reclaims, - eg. ramfs pages, shmctl(SHM_LOCK) and mlock() memory segments - 2. REFERENCED page has been referenced since last LRU list enqueue/requeue - 9. RECLAIM page will be reclaimed soon after its pageout IO completed -11. MMAP a memory mapped page -12. ANON a memory mapped page that is not part of a file -13. SWAPCACHE page is mapped to swap space, ie. has an associated swap entry -14. SWAPBACKED page is backed by swap/RAM - -The page-types tool in the tools/vm directory can be used to query the -above flags. - -Using pagemap to do something useful: - -The general procedure for using pagemap to find out about a process' memory -usage goes like this: - - 1. Read /proc/pid/maps to determine which parts of the memory space are - mapped to what. - 2. Select the maps you are interested in -- all of them, or a particular - library, or the stack or the heap, etc. - 3. Open /proc/pid/pagemap and seek to the pages you would like to examine. - 4. Read a u64 for each page from pagemap. - 5. Open /proc/kpagecount and/or /proc/kpageflags. For each PFN you just - read, seek to that entry in the file, and read the data you want. - -For example, to find the "unique set size" (USS), which is the amount of -memory that a process is using that is not shared with any other process, -you can go through every map in the process, find the PFNs, look those up -in kpagecount, and tally up the number of pages that are only referenced -once. - -Other notes: - -Reading from any of the files will return -EINVAL if you are not starting -the read on an 8-byte boundary (e.g., if you sought an odd number of bytes -into the file), or if the size of the read is not a multiple of 8 bytes. - -Before Linux 3.11 pagemap bits 55-60 were used for "page-shift" (which is -always 12 at most architectures). Since Linux 3.11 their meaning changes -after first clear of soft-dirty bits. Since Linux 4.2 they are used for -flags unconditionally. diff --git a/Documentation/vm/remap_file_pages.txt b/Documentation/vm/remap_file_pages.rst index f609142f406a1..7bef6718e3a90 100644 --- a/Documentation/vm/remap_file_pages.txt +++ b/Documentation/vm/remap_file_pages.rst @@ -1,3 +1,9 @@ +.. _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() diff --git a/Documentation/vm/slub.rst b/Documentation/vm/slub.rst new file mode 100644 index 0000000000000..3a775fd64e2db --- /dev/null +++ b/Documentation/vm/slub.rst @@ -0,0 +1,361 @@ +.. _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 name> + Enable options only for select slabs + + +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 Toggle 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. + +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 + +In case you forgot to enable debugging on the kernel command line: It is +possible to enable debugging manually when the kernel is up. Look at the +contents of:: + + /sys/kernel/slab/<slab name>/ + +Look at the writable files. Writing 1 to them will enable the +corresponding debug option. All options can be set on a slab that does +not contain objects. If the slab already contains objects then sanity checks +and tracing may only be enabled. The other options may cause the realignment +of objects. + +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 minim 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: 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 heigh + b) ``-r %d,%d`` -- specifies a range of samples to use (for example, + in ``slabinfo -X >> FOO_STATS; sleep 1;`` case, using a ``-r + 40,60`` range will plot only samples collected between 40th and + 60th seconds). + +Christoph Lameter, May 30, 2007 +Sergey Senozhatsky, October 23, 2015 diff --git a/Documentation/vm/slub.txt b/Documentation/vm/slub.txt deleted file mode 100644 index 84652419bff25..0000000000000 --- a/Documentation/vm/slub.txt +++ /dev/null @@ -1,342 +0,0 @@ -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 name> - Enable options only for select slabs - -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 Toggle 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. - -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 - -In case you forgot to enable debugging on the kernel command line: It is -possible to enable debugging manually when the kernel is up. Look at the -contents of: - -/sys/kernel/slab/<slab name>/ - -Look at the writable files. Writing 1 to them will enable the -corresponding debug option. All options can be set on a slab that does -not contain objects. If the slab already contains objects then sanity checks -and tracing may only be enabled. The other options may cause the realignment -of objects. - -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 minim 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: 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 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 heigh - b) -r %d,%d specifies a range of samples to use (for example, - in `slabinfo -X >> FOO_STATS; sleep 1;' case, using - a "-r 40,60" range will plot only samples collected - between 40th and 60th seconds). - -Christoph Lameter, May 30, 2007 -Sergey Senozhatsky, October 23, 2015 diff --git a/Documentation/vm/soft-dirty.txt b/Documentation/vm/soft-dirty.txt deleted file mode 100644 index 55684d11a1e80..0000000000000 --- a/Documentation/vm/soft-dirty.txt +++ /dev/null @@ -1,43 +0,0 @@ - SOFT-DIRTY PTEs - - The soft-dirty is a bit on a PTE which helps to track which pages a task -writes to. In order to do this tracking one should - - 1. Clear soft-dirty bits from the task's PTEs. - - This is done by writing "4" into the /proc/PID/clear_refs file of the - task in question. - - 2. Wait some time. - - 3. Read soft-dirty bits from the PTEs. - - This is done by reading from the /proc/PID/pagemap. The bit 55 of the - 64-bit qword is the soft-dirty one. If set, the respective PTE was - written to since step 1. - - - Internally, to do this tracking, the writable bit is cleared from PTEs -when the soft-dirty bit is cleared. So, after this, when the task tries to -modify a page at some virtual address the #PF occurs and the kernel sets -the soft-dirty bit on the respective PTE. - - Note, that although all the task's address space is marked as r/o after the -soft-dirty bits clear, the #PF-s that occur after that are processed fast. -This is so, since the pages are still mapped to physical memory, and thus all -the kernel does is finds this fact out and puts both writable and soft-dirty -bits on the PTE. - - While in most cases tracking memory changes by #PF-s is more than enough -there is still a scenario when we can lose soft dirty bits -- a task -unmaps a previously mapped memory region and then maps a new one at exactly -the same place. When unmap is called, the kernel internally clears PTE values -including soft dirty bits. To notify user space application about such -memory region renewal the kernel always marks new memory regions (and -expanded regions) as soft dirty. - - This feature is actively used by the checkpoint-restore project. You -can find more details about it on http://criu.org - - --- Pavel Emelyanov, Apr 9, 2013 diff --git a/Documentation/vm/split_page_table_lock b/Documentation/vm/split_page_table_lock.rst index 62842a857dab3..889b00be469f2 100644 --- a/Documentation/vm/split_page_table_lock +++ b/Documentation/vm/split_page_table_lock.rst @@ -1,3 +1,6 @@ +.. _split_page_table_lock: + +===================== Split page table lock ===================== @@ -11,6 +14,7 @@ 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; @@ -34,12 +38,13 @@ 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; @@ -47,7 +52,7 @@ Hugetlb-specific helpers: 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_page_ctor() and pgtable_page_dtor(), @@ -73,7 +78,7 @@ NOTE: pgtable_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 @@ -81,6 +86,7 @@ page of page containing the table. It shares storage with page->private 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 diff --git a/Documentation/vm/swap_numa.txt b/Documentation/vm/swap_numa.rst index d5960c9124f5f..e0466f2db8fa0 100644 --- a/Documentation/vm/swap_numa.txt +++ b/Documentation/vm/swap_numa.rst @@ -1,5 +1,8 @@ +.. _swap_numa: + +=========================================== Automatically bind swap device to numa node -------------------------------------------- +=========================================== If the system has more than one swap device and swap device has the node information, we can make use of this information to decide which swap @@ -7,15 +10,16 @@ device to use in get_swap_pages() to get better performance. How to use this feature ------------------------ +======================= Swap device has priority and that decides the order of it to be used. To make use of automatically binding, there is no need to manipulate priority settings for swap devices. e.g. on a 2 node machine, assume 2 swap devices swapA and swapB, with swapA attached to node 0 and swapB attached to node 1, are going -to be swapped on. Simply swapping them on by doing: -# swapon /dev/swapA -# swapon /dev/swapB +to be swapped on. Simply swapping them on by doing:: + + # swapon /dev/swapA + # swapon /dev/swapB Then node 0 will use the two swap devices in the order of swapA then swapB and node 1 will use the two swap devices in the order of swapB then swapA. Note @@ -24,32 +28,39 @@ that the order of them being swapped on doesn't matter. A more complex example on a 4 node machine. Assume 6 swap devices are going to be swapped on: swapA and swapB are attached to node 0, swapC is attached to node 1, swapD and swapE are attached to node 2 and swapF is attached to node3. -The way to swap them on is the same as above: -# swapon /dev/swapA -# swapon /dev/swapB -# swapon /dev/swapC -# swapon /dev/swapD -# swapon /dev/swapE -# swapon /dev/swapF - -Then node 0 will use them in the order of: -swapA/swapB -> swapC -> swapD -> swapE -> swapF +The way to swap them on is the same as above:: + + # swapon /dev/swapA + # swapon /dev/swapB + # swapon /dev/swapC + # swapon /dev/swapD + # swapon /dev/swapE + # swapon /dev/swapF + +Then node 0 will use them in the order of:: + + swapA/swapB -> swapC -> swapD -> swapE -> swapF + swapA and swapB will be used in a round robin mode before any other swap device. -node 1 will use them in the order of: -swapC -> swapA -> swapB -> swapD -> swapE -> swapF +node 1 will use them in the order of:: + + swapC -> swapA -> swapB -> swapD -> swapE -> swapF + +node 2 will use them in the order of:: + + swapD/swapE -> swapA -> swapB -> swapC -> swapF -node 2 will use them in the order of: -swapD/swapE -> swapA -> swapB -> swapC -> swapF Similaly, swapD and swapE will be used in a round robin mode before any other swap devices. -node 3 will use them in the order of: -swapF -> swapA -> swapB -> swapC -> swapD -> swapE +node 3 will use them in the order of:: + + swapF -> swapA -> swapB -> swapC -> swapD -> swapE Implementation details ----------------------- +====================== The current code uses a priority based list, swap_avail_list, to decide which swap device to use and if multiple swap devices share the same diff --git a/Documentation/vm/transhuge.rst b/Documentation/vm/transhuge.rst new file mode 100644 index 0000000000000..a8cf6809e36e4 --- /dev/null +++ b/Documentation/vm/transhuge.rst @@ -0,0 +1,197 @@ +.. _transhuge: + +============================ +Transparent Hugepage Support +============================ + +This document describes design principles Transparent Hugepage (THP) +Support and its interaction with other parts of the memory management. + +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 reference on any head/tail page would +prevent page from being split by anyone. + +.. note:: + these aren't new constraints to the GUP API, and they match the + same constrains that applies to hugetlbfs too, so any driver capable + of handling GUP on hugetlbfs will also work fine on transparent + hugepage backed mappings. + +In case you can't handle compound pages if they're returned by +follow_page, the FOLL_SPLIT bit can be specified as parameter to +follow_page, so that it will split the hugepages before returning +them. Migration for example passes FOLL_SPLIT as parameter to +follow_page because it's not hugepage aware and in fact it can't work +at all on hugetlbfs (but it instead works fine on transparent +hugepages thanks to FOLL_SPLIT). migration simply can't deal with +hugepages being returned (as it's not only checking the pfn of the +page and pinning it during the copy but it pretends to migrate the +memory in regular page sizes and with regular pte/pmd 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 +hundred if not thousand of lines of complex code to make your code +hugepage aware. + +If you're not walking pagetables but you run into a physical hugepage +but you can't handle it 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_sem in read (or write) mode to be sure an huge pmd cannot be +created from under you by khugepaged (khugepaged collapse_huge_page +takes the mmap_sem 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 to be 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 in head page's ->_refcount. + + - ->_refcount in tail pages is always zero: get_page_unless_zero() never + succeed 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 accounted 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 is optimization required to lower overhead of per-subpage mapcount +tracking. The alternative is 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 got split +for the first time, but still have PMD mapping. The additional references +go away with last compound_mapcount. + +File pages get PG_double_map set on first map of the page with PTE and +goes away when the page gets evicted from 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 page: it expects page count to be equal to +sum of mapcount of all sub-pages plus one (split_huge_page caller must +have reference for head page). + +split_huge_page uses migration entries to stabilize page->_refcount and +page->_mapcount of anonymous pages. File pages just got unmapped. + +We safe against physical memory scanners too: the only legitimate way +scanner can get 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 known 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 reference should go after split: it will stay on head page. + +Note that split_huge_pmd() doesn't have any limitation 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's also might be +counterproductive since in many cases partial unmap happens during exit(2) if +a THP crosses a VMA boundary. + +Function deferred_split_huge_page() is used to queue page for splitting. +The splitting itself will happen when we get memory pressure via shrinker +interface. diff --git a/Documentation/vm/transhuge.txt b/Documentation/vm/transhuge.txt deleted file mode 100644 index 4dde03b44ad1b..0000000000000 --- a/Documentation/vm/transhuge.txt +++ /dev/null @@ -1,527 +0,0 @@ -= Transparent Hugepage Support = - -== Objective == - -Performance critical computing applications dealing with large memory -working sets are already running on top of libhugetlbfs and in turn -hugetlbfs. Transparent Hugepage Support is an alternative means of -using huge pages for the backing of virtual memory with huge pages -that supports the automatic promotion and demotion of page sizes and -without the shortcomings of hugetlbfs. - -Currently it only works for anonymous memory mappings and tmpfs/shmem. -But in the future it can expand to other filesystems. - -The reason applications are running faster is because of two -factors. The first factor is almost completely irrelevant and it's not -of significant interest because it'll also have the downside of -requiring larger clear-page copy-page in page faults which is a -potentially negative effect. The first factor consists in taking a -single page fault for each 2M virtual region touched by userland (so -reducing the enter/exit kernel frequency by a 512 times factor). This -only matters the first time the memory is accessed for the lifetime of -a memory mapping. The second long lasting and much more important -factor will affect all subsequent accesses to the memory for the whole -runtime of the application. The second factor consist of two -components: 1) the TLB miss will run faster (especially with -virtualization using nested pagetables but almost always also on bare -metal without virtualization) and 2) a single TLB entry will be -mapping a much larger amount of virtual memory in turn reducing the -number of TLB misses. With virtualization and nested pagetables the -TLB can be mapped of larger size only if both KVM and the Linux guest -are using hugepages but a significant speedup already happens if only -one of the two is using hugepages just because of the fact the TLB -miss is going to run faster. - -== Design == - -- "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) - -Transparent Hugepage Support maximizes the usefulness of free memory -if compared to the reservation approach of hugetlbfs by allowing all -unused memory to be used as cache or other movable (or even unmovable -entities). It doesn't require reservation to prevent hugepage -allocation failures to be noticeable from userland. It allows paging -and all other advanced VM features to be available on the -hugepages. It requires no modifications for applications to take -advantage of it. - -Applications however can be further optimized to take advantage of -this feature, like for example they've been optimized before to avoid -a flood of mmap system calls for every malloc(4k). Optimizing userland -is by far not mandatory and khugepaged already can take care of long -lived page allocations even for hugepage unaware applications that -deals with large amounts of memory. - -In certain cases when hugepages are enabled system wide, application -may end up allocating more memory resources. An application may mmap a -large region but only touch 1 byte of it, in that case a 2M page might -be allocated instead of a 4k page for no good. This is why it's -possible to disable hugepages system-wide and to only have them inside -MADV_HUGEPAGE madvise regions. - -Embedded systems should enable hugepages only inside madvise regions -to eliminate any risk of wasting any precious byte of memory and to -only run faster. - -Applications that gets a lot of benefit from hugepages and that don't -risk to lose memory by using hugepages, should use -madvise(MADV_HUGEPAGE) on their critical mmapped regions. - -== sysfs == - -Transparent Hugepage Support for anonymous memory can be entirely disabled -(mostly for debugging purposes) or only enabled inside MADV_HUGEPAGE -regions (to avoid the risk of consuming more memory resources) or enabled -system wide. This can be achieved with one of: - -echo always >/sys/kernel/mm/transparent_hugepage/enabled -echo madvise >/sys/kernel/mm/transparent_hugepage/enabled -echo never >/sys/kernel/mm/transparent_hugepage/enabled - -It's also possible to limit defrag efforts in the VM to generate -anonymous hugepages in case they're not immediately free to madvise -regions or to never try to defrag memory and simply fallback to regular -pages unless hugepages are immediately available. Clearly if we spend CPU -time to defrag memory, we would expect to gain even more by the fact we -use hugepages later instead of regular pages. This isn't always -guaranteed, but it may be more likely in case the allocation is for a -MADV_HUGEPAGE region. - -echo always >/sys/kernel/mm/transparent_hugepage/defrag -echo defer >/sys/kernel/mm/transparent_hugepage/defrag -echo defer+madvise >/sys/kernel/mm/transparent_hugepage/defrag -echo madvise >/sys/kernel/mm/transparent_hugepage/defrag -echo never >/sys/kernel/mm/transparent_hugepage/defrag - -"always" means that an application requesting THP will stall on allocation -failure and directly reclaim pages and compact memory in an effort to -allocate a THP immediately. This may be desirable for virtual machines -that benefit heavily from THP use and are willing to delay the VM start -to utilise them. - -"defer" means that an application will wake kswapd in the background -to reclaim pages and wake kcompactd to compact memory so that THP is -available in the near future. It's the responsibility of khugepaged -to then install the THP pages later. - -"defer+madvise" will enter direct reclaim and compaction like "always", but -only for regions that have used madvise(MADV_HUGEPAGE); all other regions -will wake kswapd in the background to reclaim pages and wake kcompactd to -compact memory so that THP is available in the near future. - -"madvise" will enter direct reclaim like "always" but only for regions -that are have used madvise(MADV_HUGEPAGE). This is the default behaviour. - -"never" should be self-explanatory. - -By default kernel tries to use huge zero page on read page fault to -anonymous mapping. It's possible to disable huge zero page by writing 0 -or enable it back by writing 1: - -echo 0 >/sys/kernel/mm/transparent_hugepage/use_zero_page -echo 1 >/sys/kernel/mm/transparent_hugepage/use_zero_page - -Some userspace (such as a test program, or an optimized memory allocation -library) may want to know the size (in bytes) of a transparent hugepage: - -cat /sys/kernel/mm/transparent_hugepage/hpage_pmd_size - -khugepaged will be automatically started when -transparent_hugepage/enabled is set to "always" or "madvise, and it'll -be automatically shutdown if it's set to "never". - -khugepaged runs usually at low frequency so while one may not want to -invoke defrag algorithms synchronously during the page faults, it -should be worth invoking defrag at least in khugepaged. However it's -also possible to disable defrag in khugepaged by writing 0 or enable -defrag in khugepaged by writing 1: - -echo 0 >/sys/kernel/mm/transparent_hugepage/khugepaged/defrag -echo 1 >/sys/kernel/mm/transparent_hugepage/khugepaged/defrag - -You can also control how many pages khugepaged should scan at each -pass: - -/sys/kernel/mm/transparent_hugepage/khugepaged/pages_to_scan - -and how many milliseconds to wait in khugepaged between each pass (you -can set this to 0 to run khugepaged at 100% utilization of one core): - -/sys/kernel/mm/transparent_hugepage/khugepaged/scan_sleep_millisecs - -and how many milliseconds to wait in khugepaged if there's an hugepage -allocation failure to throttle the next allocation attempt. - -/sys/kernel/mm/transparent_hugepage/khugepaged/alloc_sleep_millisecs - -The khugepaged progress can be seen in the number of pages collapsed: - -/sys/kernel/mm/transparent_hugepage/khugepaged/pages_collapsed - -for each pass: - -/sys/kernel/mm/transparent_hugepage/khugepaged/full_scans - -max_ptes_none specifies how many extra small pages (that are -not already mapped) can be allocated when collapsing a group -of small pages into one large page. - -/sys/kernel/mm/transparent_hugepage/khugepaged/max_ptes_none - -A higher value leads to use additional memory for programs. -A lower value leads to gain less thp performance. Value of -max_ptes_none can waste cpu time very little, you can -ignore it. - -max_ptes_swap specifies how many pages can be brought in from -swap when collapsing a group of pages into a transparent huge page. - -/sys/kernel/mm/transparent_hugepage/khugepaged/max_ptes_swap - -A higher value can cause excessive swap IO and waste -memory. A lower value can prevent THPs from being -collapsed, resulting fewer pages being collapsed into -THPs, and lower memory access performance. - -== Boot parameter == - -You can change the sysfs boot time defaults of Transparent Hugepage -Support by passing the parameter "transparent_hugepage=always" or -"transparent_hugepage=madvise" or "transparent_hugepage=never" -(without "") to the kernel command line. - -== Hugepages in tmpfs/shmem == - -You can control hugepage allocation policy in tmpfs with mount option -"huge=". It can have following values: - - - "always": - Attempt to allocate huge pages every time we need a new page; - - - "never": - Do not allocate huge pages; - - - "within_size": - Only allocate huge page if it will be fully within i_size. - Also respect fadvise()/madvise() hints; - - - "advise: - Only allocate huge pages if requested with fadvise()/madvise(); - -The default policy is "never". - -"mount -o remount,huge= /mountpoint" works fine after mount: remounting -huge=never will not attempt to break up huge pages at all, just stop more -from being allocated. - -There's also sysfs knob to control hugepage allocation policy for internal -shmem mount: /sys/kernel/mm/transparent_hugepage/shmem_enabled. The mount -is used for SysV SHM, memfds, shared anonymous mmaps (of /dev/zero or -MAP_ANONYMOUS), GPU drivers' DRM objects, Ashmem. - -In addition to policies listed above, shmem_enabled allows two further -values: - - - "deny": - For use in emergencies, to force the huge option off from - all mounts; - - "force": - Force the huge option on for all - very useful for testing; - -== Need of application restart == - -The transparent_hugepage/enabled values and tmpfs mount option only affect -future behavior. So to make them effective you need to restart any -application that could have been using hugepages. This also applies to the -regions registered in khugepaged. - -== Monitoring usage == - -The number of anonymous transparent huge pages currently used by the -system is available by reading the AnonHugePages field in /proc/meminfo. -To identify what applications are using anonymous transparent huge pages, -it is necessary to read /proc/PID/smaps and count the AnonHugePages fields -for each mapping. - -The number of file transparent huge pages mapped to userspace is available -by reading ShmemPmdMapped and ShmemHugePages fields in /proc/meminfo. -To identify what applications are mapping file transparent huge pages, it -is necessary to read /proc/PID/smaps and count the FileHugeMapped fields -for each mapping. - -Note that reading the smaps file is expensive and reading it -frequently will incur overhead. - -There are a number of counters in /proc/vmstat that may be used to -monitor how successfully the system is providing huge pages for use. - -thp_fault_alloc is incremented every time a huge page is successfully - allocated to handle a page fault. This applies to both the - first time a page is faulted and for COW faults. - -thp_collapse_alloc is incremented by khugepaged when it has found - a range of pages to collapse into one huge page and has - successfully allocated a new huge page to store the data. - -thp_fault_fallback is incremented if a page fault fails to allocate - a huge page and instead falls back to using small pages. - -thp_collapse_alloc_failed is incremented if khugepaged found a range - of pages that should be collapsed into one huge page but failed - the allocation. - -thp_file_alloc is incremented every time a file huge page is successfully - allocated. - -thp_file_mapped is incremented every time a file huge page is mapped into - user address space. - -thp_split_page is incremented every time a huge page is split into base - pages. This can happen for a variety of reasons but a common - reason is that a huge page is old and is being reclaimed. - This action implies splitting all PMD the page mapped with. - -thp_split_page_failed is incremented if kernel fails to split huge - page. This can happen if the page was pinned by somebody. - -thp_deferred_split_page is incremented when a huge page is put onto split - queue. This happens when a huge page is partially unmapped and - splitting it would free up some memory. Pages on split queue are - going to be split under memory pressure. - -thp_split_pmd is incremented every time a PMD split into table of PTEs. - This can happen, for instance, when application calls mprotect() or - munmap() on part of huge page. It doesn't split huge page, only - page table entry. - -thp_zero_page_alloc is incremented every time a huge zero page is - successfully allocated. It includes allocations which where - dropped due race with other allocation. Note, it doesn't count - every map of the huge zero page, only its allocation. - -thp_zero_page_alloc_failed is incremented if kernel fails to allocate - huge zero page and falls back to using small pages. - -As the system ages, allocating huge pages may be expensive as the -system uses memory compaction to copy data around memory to free a -huge page for use. There are some counters in /proc/vmstat to help -monitor this overhead. - -compact_stall is incremented every time a process stalls to run - memory compaction so that a huge page is free for use. - -compact_success is incremented if the system compacted memory and - freed a huge page for use. - -compact_fail is incremented if the system tries to compact memory - but failed. - -compact_pages_moved is incremented each time a page is moved. If - this value is increasing rapidly, it implies that the system - is copying a lot of data to satisfy the huge page allocation. - It is possible that the cost of copying exceeds any savings - from reduced TLB misses. - -compact_pagemigrate_failed is incremented when the underlying mechanism - for moving a page failed. - -compact_blocks_moved is incremented each time memory compaction examines - a huge page aligned range of pages. - -It is possible to establish how long the stalls were using the function -tracer to record how long was spent in __alloc_pages_nodemask and -using the mm_page_alloc tracepoint to identify which allocations were -for huge pages. - -== 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 reference on any head/tail page would -prevent page from being split by anyone. - -NOTE: these aren't new constraints to the GUP API, and they match the -same constrains that applies to hugetlbfs too, so any driver capable -of handling GUP on hugetlbfs will also work fine on transparent -hugepage backed mappings. - -In case you can't handle compound pages if they're returned by -follow_page, the FOLL_SPLIT bit can be specified as parameter to -follow_page, so that it will split the hugepages before returning -them. Migration for example passes FOLL_SPLIT as parameter to -follow_page because it's not hugepage aware and in fact it can't work -at all on hugetlbfs (but it instead works fine on transparent -hugepages thanks to FOLL_SPLIT). migration simply can't deal with -hugepages being returned (as it's not only checking the pfn of the -page and pinning it during the copy but it pretends to migrate the -memory in regular page sizes and with regular pte/pmd mappings). - -== Optimizing the applications == - -To be guaranteed that the kernel will map a 2M page immediately in any -memory region, the mmap region has to be hugepage naturally -aligned. posix_memalign() can provide that guarantee. - -== Hugetlbfs == - -You can use hugetlbfs on a kernel that has transparent hugepage -support enabled just fine as always. No difference can be noted in -hugetlbfs other than there will be less overall fragmentation. All -usual features belonging to hugetlbfs are preserved and -unaffected. libhugetlbfs will also work fine as usual. - -== 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 -hundred if not thousand of lines of complex code to make your code -hugepage aware. - -If you're not walking pagetables but you run into a physical hugepage -but you can't handle it 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_sem in read (or write) mode to be sure an huge pmd cannot be -created from under you by khugepaged (khugepaged collapse_huge_page -takes the mmap_sem 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 to be 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 in head page's ->_refcount. - - - ->_refcount in tail pages is always zero: get_page_unless_zero() never - succeed 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 accounted 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 is optimization required to lower overhead of per-subpage mapcount -tracking. The alternative is 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 got split -for the first time, but still have PMD mapping. The additional references -go away with last compound_mapcount. - -File pages get PG_double_map set on first map of the page with PTE and -goes away when the page gets evicted from 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 page: it expects page count to be equal to -sum of mapcount of all sub-pages plus one (split_huge_page caller must -have reference for head page). - -split_huge_page uses migration entries to stabilize page->_refcount and -page->_mapcount of anonymous pages. File pages just got unmapped. - -We safe against physical memory scanners too: the only legitimate way -scanner can get 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 known 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 reference should go after split: it will stay on head page. - -Note that split_huge_pmd() doesn't have any limitation 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's also might be -counterproductive since in many cases partial unmap happens during exit(2) if -a THP crosses a VMA boundary. - -Function deferred_split_huge_page() is used to queue page for splitting. -The splitting itself will happen when we get memory pressure via shrinker -interface. diff --git a/Documentation/vm/unevictable-lru.txt b/Documentation/vm/unevictable-lru.rst index e147185724766..fdd84cb8d511f 100644 --- a/Documentation/vm/unevictable-lru.txt +++ b/Documentation/vm/unevictable-lru.rst @@ -1,37 +1,13 @@ - ============================== - UNEVICTABLE LRU INFRASTRUCTURE - ============================== - -======== -CONTENTS -======== - - (*) The Unevictable LRU - - - The unevictable page list. - - Memory control group interaction. - - Marking address spaces unevictable. - - Detecting Unevictable Pages. - - vmscan's handling of unevictable pages. - - (*) mlock()'d pages. - - - History. - - Basic management. - - mlock()/mlockall() system call handling. - - Filtering special vmas. - - munlock()/munlockall() system call handling. - - Migrating mlocked pages. - - Compacting mlocked pages. - - mmap(MAP_LOCKED) system call handling. - - munmap()/exit()/exec() system call handling. - - try_to_unmap(). - - try_to_munlock() reverse map scan. - - Page reclaim in shrink_*_list(). +.. _unevictable_lru: +============================== +Unevictable LRU Infrastructure +============================== -============ -INTRODUCTION +.. contents:: :local: + + +Introduction ============ This document describes the Linux memory manager's "Unevictable LRU" @@ -46,8 +22,8 @@ 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 =================== The Unevictable LRU facility adds an additional LRU list to track unevictable @@ -66,17 +42,17 @@ completely unresponsive. The unevictable list addresses the following classes of unevictable pages: - (*) Those owned by ramfs. + * Those owned by ramfs. - (*) Those mapped into SHM_LOCK'd shared memory regions. + * Those mapped into SHM_LOCK'd shared memory regions. - (*) Those mapped into VM_LOCKED [mlock()ed] VMAs. + * Those mapped into VM_LOCKED [mlock()ed] VMAs. The infrastructure may also be able to handle other conditions that make pages unevictable, either by definition or by circumstance, in the future. -THE UNEVICTABLE PAGE LIST +The Unevictable Page List ------------------------- The Unevictable LRU infrastructure consists of an additional, per-zone, LRU list @@ -118,7 +94,7 @@ the unevictable list when one task has the page isolated from the LRU and other tasks are changing the "evictability" state of the page. -MEMORY CONTROL GROUP INTERACTION +Memory Control Group Interaction -------------------------------- The unevictable LRU facility interacts with the memory control group [aka @@ -144,7 +120,9 @@ effects: the control group to thrash or to OOM-kill tasks. -MARKING ADDRESS SPACES UNEVICTABLE +.. _mark_addr_space_unevict: + +Marking Address Spaces Unevictable ---------------------------------- For facilities such as ramfs none of the pages attached to the address space @@ -152,15 +130,15 @@ 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); + * ``void mapping_set_unevictable(struct address_space *mapping);`` Mark the address space as being completely unevictable. - (*) void mapping_clear_unevictable(struct address_space *mapping); + * ``void mapping_clear_unevictable(struct address_space *mapping);`` Mark the address space as being evictable. - (*) int mapping_unevictable(struct address_space *mapping); + * ``int mapping_unevictable(struct address_space *mapping);`` Query the address space, and return true if it is completely unevictable. @@ -177,12 +155,13 @@ These are currently used in two places in the kernel: ensure they're in memory. -DETECTING UNEVICTABLE PAGES +Detecting Unevictable Pages --------------------------- The function page_evictable() in vmscan.c determines whether a page is -evictable or not using the query function outlined above [see section "Marking -address spaces unevictable"] to check the AS_UNEVICTABLE flag. +evictable or not using the query function outlined above [see section +:ref:`Marking address spaces unevictable <mark_addr_space_unevict>`] +to check the AS_UNEVICTABLE flag. For address spaces that are so marked after being populated (as SHM regions might be), the lock action (eg: SHM_LOCK) can be lazy, and need not populate @@ -202,7 +181,7 @@ 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 +Vmscan's Handling of Unevictable Pages -------------------------------------- If unevictable pages are culled in the fault path, or moved to the unevictable @@ -233,8 +212,7 @@ extra evictabilty checks should not occur in the majority of calls to putback_lru_page(). -============= -MLOCKED PAGES +MLOCKED Pages ============= The unevictable page list is also useful for mlock(), in addition to ramfs and @@ -242,7 +220,7 @@ SYSV SHM. Note that mlock() is only available in CONFIG_MMU=y situations; in NOMMU situations, all mappings are effectively mlocked. -HISTORY +History ------- The "Unevictable mlocked Pages" infrastructure is based on work originally @@ -263,7 +241,7 @@ replaced by walking the reverse map to determine whether any VM_LOCKED VMAs mapped the page. More on this below. -BASIC MANAGEMENT +Basic Management ---------------- mlocked pages - pages mapped into a VM_LOCKED VMA - are a class of unevictable @@ -304,10 +282,10 @@ mlocked pages become unlocked and rescued from the unevictable list when: (4) before a page is COW'd in a VM_LOCKED VMA. -mlock()/mlockall() SYSTEM CALL HANDLING +mlock()/mlockall() System Call Handling --------------------------------------- -Both [do_]mlock() and [do_]mlockall() system call handlers call mlock_fixup() +Both [do\_]mlock() and [do\_]mlockall() system call handlers call mlock_fixup() for each VMA in the range specified by the call. In the case of mlockall(), this is the entire active address space of the task. Note that mlock_fixup() is used for both mlocking and munlocking a range of memory. A call to mlock() @@ -351,7 +329,7 @@ mlock_vma_page() is unable to isolate the page from the LRU, vmscan will handle it later if and when it attempts to reclaim the page. -FILTERING SPECIAL VMAS +Filtering Special VMAs ---------------------- mlock_fixup() filters several classes of "special" VMAs: @@ -379,8 +357,9 @@ VM_LOCKED flag. Therefore, we won't have to deal with them later during munlock(), munmap() or task exit. Neither does mlock_fixup() account these VMAs against the task's "locked_vm". +.. _munlock_munlockall_handling: -munlock()/munlockall() SYSTEM CALL HANDLING +munlock()/munlockall() System Call Handling ------------------------------------------- The munlock() and munlockall() system calls are handled by the same functions - @@ -426,7 +405,7 @@ This is fine, because we'll catch it later if and if vmscan tries to reclaim the page. This should be relatively rare. -MIGRATING MLOCKED PAGES +Migrating MLOCKED Pages ----------------------- A page that is being migrated has been isolated from the LRU lists and is held @@ -451,7 +430,7 @@ list because of a race between munlock and migration, page migration uses the putback_lru_page() function to add migrated pages back to the LRU. -COMPACTING MLOCKED PAGES +Compacting MLOCKED Pages ------------------------ The unevictable LRU can be scanned for compactable regions and the default @@ -461,7 +440,7 @@ unevictable LRU is enabled, the work of compaction is mostly handled by the page migration code and the same work flow as described in MIGRATING MLOCKED PAGES will apply. -MLOCKING TRANSPARENT HUGE PAGES +MLOCKING Transparent Huge Pages ------------------------------- A transparent huge page is represented by a single entry on an LRU list. @@ -483,7 +462,7 @@ to unevictable LRU and the rest can be reclaimed. See also comment in follow_trans_huge_pmd(). -mmap(MAP_LOCKED) SYSTEM CALL HANDLING +mmap(MAP_LOCKED) System Call Handling ------------------------------------- In addition the mlock()/mlockall() system calls, an application can request @@ -514,7 +493,7 @@ memory range accounted as locked_vm, as the protections could be changed later and pages allocated into that region. -munmap()/exit()/exec() SYSTEM CALL HANDLING +munmap()/exit()/exec() System Call Handling ------------------------------------------- When unmapping an mlocked region of memory, whether by an explicit call to @@ -568,16 +547,18 @@ munlock or munmap system calls, mm teardown (munlock_vma_pages_all), reclaim, holepunching, and truncation of file pages and their anonymous COWed pages. -try_to_munlock() REVERSE MAP SCAN +try_to_munlock() Reverse Map Scan --------------------------------- - [!] TODO/FIXME: a better name might be page_mlocked() - analogous to the - page_referenced() reverse map walker. +.. warning:: + [!] TODO/FIXME: a better name might be page_mlocked() - analogous to the + page_referenced() reverse map walker. -When munlock_vma_page() [see section "munlock()/munlockall() System Call -Handling" above] tries to munlock a page, it needs to determine whether or not -the page is mapped by any VM_LOCKED VMA without actually attempting to unmap -all PTEs from the page. For this purpose, the unevictable/mlock infrastructure +When munlock_vma_page() [see section :ref:`munlock()/munlockall() System Call +Handling <munlock_munlockall_handling>` above] tries to munlock a +page, it needs to determine whether or not the page is mapped by any +VM_LOCKED VMA without actually attempting to unmap all PTEs from the +page. For this purpose, the unevictable/mlock infrastructure introduced a variant of try_to_unmap() called try_to_munlock(). try_to_munlock() calls the same functions as try_to_unmap() for anonymous and @@ -595,7 +576,7 @@ large region or tearing down a large address space that has been mlocked via mlockall(), overall this is a fairly rare event. -PAGE RECLAIM IN shrink_*_list() +Page Reclaim in shrink_*_list() ------------------------------- shrink_active_list() culls any obviously unevictable pages - i.e. diff --git a/Documentation/vm/userfaultfd.txt b/Documentation/vm/userfaultfd.txt deleted file mode 100644 index bb2f945f87ab6..0000000000000 --- a/Documentation/vm/userfaultfd.txt +++ /dev/null @@ -1,229 +0,0 @@ -= Userfaultfd = - -== Objective == - -Userfaults allow the implementation of on-demand paging from userland -and more generally they allow userland to take control of various -memory page faults, something otherwise only the kernel code could do. - -For example userfaults allows a proper and more optimal implementation -of the PROT_NONE+SIGSEGV trick. - -== Design == - -Userfaults are delivered and resolved through the userfaultfd syscall. - -The userfaultfd (aside from registering and unregistering virtual -memory ranges) provides two primary functionalities: - -1) read/POLLIN protocol to notify a userland thread of the faults - happening - -2) various UFFDIO_* ioctls that can manage the virtual memory regions - registered in the userfaultfd that allows userland to efficiently - resolve the userfaults it receives via 1) or to manage the virtual - memory in the background - -The real advantage of userfaults if compared to regular virtual memory -management of mremap/mprotect is that the userfaults in all their -operations never involve heavyweight structures like vmas (in fact the -userfaultfd runtime load never takes the mmap_sem for writing). - -Vmas are not suitable for page- (or hugepage) granular fault tracking -when dealing with virtual address spaces that could span -Terabytes. Too many vmas would be needed for that. - -The userfaultfd once opened by invoking the syscall, can also be -passed using unix domain sockets to a manager process, so the same -manager process could handle the userfaults of a multitude of -different processes without them being aware about what is going on -(well of course unless they later try to use the userfaultfd -themselves on the same region the manager is already tracking, which -is a corner case that would currently return -EBUSY). - -== API == - -When first opened the userfaultfd must be enabled invoking the -UFFDIO_API ioctl specifying a uffdio_api.api value set to UFFD_API (or -a later API version) which will specify the read/POLLIN protocol -userland intends to speak on the UFFD and the uffdio_api.features -userland requires. The UFFDIO_API ioctl if successful (i.e. if the -requested uffdio_api.api is spoken also by the running kernel and the -requested features are going to be enabled) will return into -uffdio_api.features and uffdio_api.ioctls two 64bit bitmasks of -respectively all the available features of the read(2) protocol and -the generic ioctl available. - -The uffdio_api.features bitmask returned by the UFFDIO_API ioctl -defines what memory types are supported by the userfaultfd and what -events, except page fault notifications, may be generated. - -If the kernel supports registering userfaultfd ranges on hugetlbfs -virtual memory areas, UFFD_FEATURE_MISSING_HUGETLBFS will be set in -uffdio_api.features. Similarly, UFFD_FEATURE_MISSING_SHMEM will be -set if the kernel supports registering userfaultfd ranges on shared -memory (covering all shmem APIs, i.e. tmpfs, IPCSHM, /dev/zero -MAP_SHARED, memfd_create, etc). - -The userland application that wants to use userfaultfd with hugetlbfs -or shared memory need to set the corresponding flag in -uffdio_api.features to enable those features. - -If the userland desires to receive notifications for events other than -page faults, it has to verify that uffdio_api.features has appropriate -UFFD_FEATURE_EVENT_* bits set. These events are described in more -detail below in "Non-cooperative userfaultfd" section. - -Once the userfaultfd has been enabled the UFFDIO_REGISTER ioctl should -be invoked (if present in the returned uffdio_api.ioctls bitmask) to -register a memory range in the userfaultfd by setting the -uffdio_register structure accordingly. The uffdio_register.mode -bitmask will specify to the kernel which kind of faults to track for -the range (UFFDIO_REGISTER_MODE_MISSING would track missing -pages). The UFFDIO_REGISTER ioctl will return the -uffdio_register.ioctls bitmask of ioctls that are suitable to resolve -userfaults on the range registered. Not all ioctls will necessarily be -supported for all memory types depending on the underlying virtual -memory backend (anonymous memory vs tmpfs vs real filebacked -mappings). - -Userland can use the uffdio_register.ioctls to manage the virtual -address space in the background (to add or potentially also remove -memory from the userfaultfd registered range). This means a userfault -could be triggering just before userland maps in the background the -user-faulted page. - -The primary ioctl to resolve userfaults is UFFDIO_COPY. That -atomically copies a page into the userfault registered range and wakes -up the blocked userfaults (unless uffdio_copy.mode & -UFFDIO_COPY_MODE_DONTWAKE is set). Other ioctl works similarly to -UFFDIO_COPY. They're atomic as in guaranteeing that nothing can see an -half copied page since it'll keep userfaulting until the copy has -finished. - -== QEMU/KVM == - -QEMU/KVM is using the userfaultfd syscall to implement postcopy live -migration. Postcopy live migration is one form of memory -externalization consisting of a virtual machine running with part or -all of its memory residing on a different node in the cloud. The -userfaultfd abstraction is generic enough that not a single line of -KVM kernel code had to be modified in order to add postcopy live -migration to QEMU. - -Guest async page faults, FOLL_NOWAIT and all other GUP features work -just fine in combination with userfaults. Userfaults trigger async -page faults in the guest scheduler so those guest processes that -aren't waiting for userfaults (i.e. network bound) can keep running in -the guest vcpus. - -It is generally beneficial to run one pass of precopy live migration -just before starting postcopy live migration, in order to avoid -generating userfaults for readonly guest regions. - -The implementation of postcopy live migration currently uses one -single bidirectional socket but in the future two different sockets -will be used (to reduce the latency of the userfaults to the minimum -possible without having to decrease /proc/sys/net/ipv4/tcp_wmem). - -The QEMU in the source node writes all pages that it knows are missing -in the destination node, into the socket, and the migration thread of -the QEMU running in the destination node runs UFFDIO_COPY|ZEROPAGE -ioctls on the userfaultfd in order to map the received pages into the -guest (UFFDIO_ZEROCOPY is used if the source page was a zero page). - -A different postcopy thread in the destination node listens with -poll() to the userfaultfd in parallel. When a POLLIN event is -generated after a userfault triggers, the postcopy thread read() from -the userfaultfd and receives the fault address (or -EAGAIN in case the -userfault was already resolved and waken by a UFFDIO_COPY|ZEROPAGE run -by the parallel QEMU migration thread). - -After the QEMU postcopy thread (running in the destination node) gets -the userfault address it writes the information about the missing page -into the socket. The QEMU source node receives the information and -roughly "seeks" to that page address and continues sending all -remaining missing pages from that new page offset. Soon after that -(just the time to flush the tcp_wmem queue through the network) the -migration thread in the QEMU running in the destination node will -receive the page that triggered the userfault and it'll map it as -usual with the UFFDIO_COPY|ZEROPAGE (without actually knowing if it -was spontaneously sent by the source or if it was an urgent page -requested through a userfault). - -By the time the userfaults start, the QEMU in the destination node -doesn't need to keep any per-page state bitmap relative to the live -migration around and a single per-page bitmap has to be maintained in -the QEMU running in the source node to know which pages are still -missing in the destination node. The bitmap in the source node is -checked to find which missing pages to send in round robin and we seek -over it when receiving incoming userfaults. After sending each page of -course the bitmap is updated accordingly. It's also useful to avoid -sending the same page twice (in case the userfault is read by the -postcopy thread just before UFFDIO_COPY|ZEROPAGE runs in the migration -thread). - -== Non-cooperative userfaultfd == - -When the userfaultfd is monitored by an external manager, the manager -must be able to track changes in the process virtual memory -layout. Userfaultfd can notify the manager about such changes using -the same read(2) protocol as for the page fault notifications. The -manager has to explicitly enable these events by setting appropriate -bits in uffdio_api.features passed to UFFDIO_API ioctl: - -UFFD_FEATURE_EVENT_FORK - enable userfaultfd hooks for fork(). When -this feature is enabled, the userfaultfd context of the parent process -is duplicated into the newly created process. The manager receives -UFFD_EVENT_FORK with file descriptor of the new userfaultfd context in -the uffd_msg.fork. - -UFFD_FEATURE_EVENT_REMAP - enable notifications about mremap() -calls. When the non-cooperative process moves a virtual memory area to -a different location, the manager will receive UFFD_EVENT_REMAP. The -uffd_msg.remap will contain the old and new addresses of the area and -its original length. - -UFFD_FEATURE_EVENT_REMOVE - enable notifications about -madvise(MADV_REMOVE) and madvise(MADV_DONTNEED) calls. The event -UFFD_EVENT_REMOVE will be generated upon these calls to madvise. The -uffd_msg.remove will contain start and end addresses of the removed -area. - -UFFD_FEATURE_EVENT_UNMAP - enable notifications about memory -unmapping. The manager will get UFFD_EVENT_UNMAP with uffd_msg.remove -containing start and end addresses of the unmapped area. - -Although the UFFD_FEATURE_EVENT_REMOVE and UFFD_FEATURE_EVENT_UNMAP -are pretty similar, they quite differ in the action expected from the -userfaultfd manager. In the former case, the virtual memory is -removed, but the area is not, the area remains monitored by the -userfaultfd, and if a page fault occurs in that area it will be -delivered to the manager. The proper resolution for such page fault is -to zeromap the faulting address. However, in the latter case, when an -area is unmapped, either explicitly (with munmap() system call), or -implicitly (e.g. during mremap()), the area is removed and in turn the -userfaultfd context for such area disappears too and the manager will -not get further userland page faults from the removed area. Still, the -notification is required in order to prevent manager from using -UFFDIO_COPY on the unmapped area. - -Unlike userland page faults which have to be synchronous and require -explicit or implicit wakeup, all the events are delivered -asynchronously and the non-cooperative process resumes execution as -soon as manager executes read(). The userfaultfd manager should -carefully synchronize calls to UFFDIO_COPY with the events -processing. To aid the synchronization, the UFFDIO_COPY ioctl will -return -ENOSPC when the monitored process exits at the time of -UFFDIO_COPY, and -ENOENT, when the non-cooperative process has changed -its virtual memory layout simultaneously with outstanding UFFDIO_COPY -operation. - -The current asynchronous model of the event delivery is optimal for -single threaded non-cooperative userfaultfd manager implementations. A -synchronous event delivery model can be added later as a new -userfaultfd feature to facilitate multithreading enhancements of the -non cooperative manager, for example to allow UFFDIO_COPY ioctls to -run in parallel to the event reception. Single threaded -implementations should continue to use the current async event -delivery model instead. diff --git a/Documentation/vm/z3fold.txt b/Documentation/vm/z3fold.rst index 38e4dac810b62..224e3c61d6863 100644 --- a/Documentation/vm/z3fold.txt +++ b/Documentation/vm/z3fold.rst @@ -1,5 +1,8 @@ +.. _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. @@ -7,6 +10,7 @@ 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 diff --git a/Documentation/vm/zsmalloc.txt b/Documentation/vm/zsmalloc.rst index 64ed63c4f69d6..6e79893d61326 100644 --- a/Documentation/vm/zsmalloc.txt +++ b/Documentation/vm/zsmalloc.rst @@ -1,5 +1,8 @@ +.. _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 @@ -31,40 +34,49 @@ 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: +``/sys/kernel/debug/zsmalloc/<user name>``. Here is a sample of stat output:: -# cat /sys/kernel/debug/zsmalloc/zram0/classes + # 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 -class: index -size: object size zspage stores -almost_empty: the number of ZS_ALMOST_EMPTY zspages(see below) -almost_full: the number of ZS_ALMOST_FULL zspages(see below) -obj_allocated: the number of objects allocated -obj_used: the number of objects allocated to the user -pages_used: the number of pages allocated for the class -pages_per_zspage: the number of 0-order pages to make a zspage +We assign a zspage to ZS_ALMOST_EMPTY fullness group when n <= N / f, where -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) +* 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 + +* ZS_ALMOST_FULL when n > N / f +* ZS_EMPTY when n == 0 +* ZS_FULL when n == N diff --git a/Documentation/vm/zswap.txt b/Documentation/vm/zswap.rst index 0b3a1148f9f04..1444ecd40911a 100644 --- a/Documentation/vm/zswap.txt +++ b/Documentation/vm/zswap.rst @@ -1,4 +1,11 @@ -Overview: +.. _zswap: + +===== +zswap +===== + +Overview +======== Zswap is a lightweight compressed cache for swap pages. It takes pages that are in the process of being swapped out and attempts to compress them into a @@ -7,32 +14,34 @@ for potentially reduced swap I/O. This trade-off can also result in a significant performance improvement if reads from the compressed cache are faster than reads from a swap device. -NOTE: Zswap is a new feature as of v3.11 and interacts heavily with memory -reclaim. This interaction has not been fully explored on the large set of -potential configurations and workloads that exist. For this reason, zswap -is a work in progress and should be considered experimental. +.. note:: + Zswap is a new feature as of v3.11 and interacts heavily with memory + reclaim. This interaction has not been fully explored on the large set of + potential configurations and workloads that exist. For this reason, zswap + is a work in progress and should be considered experimental. + + Some potential benefits: -Some potential benefits: * Desktop/laptop users with limited RAM capacities can mitigate the - performance impact of swapping. + performance impact of swapping. * Overcommitted guests that share a common I/O resource can - dramatically reduce their swap I/O pressure, avoiding heavy handed I/O - throttling by the hypervisor. This allows more work to get done with less - impact to the guest workload and guests sharing the I/O subsystem + dramatically reduce their swap I/O pressure, avoiding heavy handed I/O + throttling by the hypervisor. This allows more work to get done with less + impact to the guest workload and guests sharing the I/O subsystem * Users with SSDs as swap devices can extend the life of the device by - drastically reducing life-shortening writes. + drastically reducing life-shortening writes. Zswap evicts pages from compressed cache on an LRU basis to the backing swap device when the compressed pool reaches its size limit. This requirement had been identified in prior community discussions. Zswap is disabled by default but can be enabled at boot time by setting -the "enabled" attribute to 1 at boot time. ie: zswap.enabled=1. Zswap +the ``enabled`` attribute to 1 at boot time. ie: ``zswap.enabled=1``. Zswap can also be enabled and disabled at runtime using the sysfs interface. An example command to enable zswap at runtime, assuming sysfs is mounted -at /sys, is: +at ``/sys``, is:: -echo 1 > /sys/module/zswap/parameters/enabled + echo 1 > /sys/module/zswap/parameters/enabled When zswap is disabled at runtime it will stop storing pages that are being swapped out. However, it will _not_ immediately write out or fault @@ -43,7 +52,8 @@ pages out of the compressed pool, a swapoff on the swap device(s) will fault back into memory all swapped out pages, including those in the compressed pool. -Design: +Design +====== Zswap receives pages for compression through the Frontswap API and is able to evict pages from its own compressed pool on an LRU basis and write them back to @@ -53,12 +63,12 @@ Zswap makes use of zpool for the managing the compressed memory pool. Each allocation in zpool is not directly accessible by address. Rather, a handle is returned by the allocation routine and that handle must be mapped before being accessed. The compressed memory pool grows on demand and shrinks as compressed -pages are freed. The pool is not preallocated. By default, a zpool of type -zbud is created, but it can be selected at boot time by setting the "zpool" -attribute, e.g. zswap.zpool=zbud. It can also be changed at runtime using the -sysfs "zpool" attribute, e.g. +pages are freed. The pool is not preallocated. By default, a zpool +of type zbud is created, but it can be selected at boot time by +setting the ``zpool`` attribute, e.g. ``zswap.zpool=zbud``. It can +also be changed at runtime using the sysfs ``zpool`` attribute, e.g.:: -echo zbud > /sys/module/zswap/parameters/zpool + echo zbud > /sys/module/zswap/parameters/zpool The zbud type zpool allocates exactly 1 page to store 2 compressed pages, which means the compression ratio will always be 2:1 or worse (because of half-full @@ -83,14 +93,16 @@ via frontswap, to free the compressed entry. Zswap seeks to be simple in its policies. Sysfs attributes allow for one user controlled policy: + * max_pool_percent - The maximum percentage of memory that the compressed - pool can occupy. + pool can occupy. -The default compressor is lzo, but it can be selected at boot time by setting -the “compressor” attribute, e.g. zswap.compressor=lzo. It can also be changed -at runtime using the sysfs "compressor" attribute, e.g. +The default compressor is lzo, but it can be selected at boot time by +setting the ``compressor`` attribute, e.g. ``zswap.compressor=lzo``. +It can also be changed at runtime using the sysfs "compressor" +attribute, e.g.:: -echo lzo > /sys/module/zswap/parameters/compressor + echo lzo > /sys/module/zswap/parameters/compressor When the zpool and/or compressor parameter is changed at runtime, any existing compressed pages are not modified; they are left in their own zpool. When a @@ -106,11 +118,12 @@ compressed length of the page is set to zero and the pattern or same-filled value is stored. Same-value filled pages identification feature is enabled by default and can be -disabled at boot time by setting the "same_filled_pages_enabled" attribute to 0, -e.g. zswap.same_filled_pages_enabled=0. It can also be enabled and disabled at -runtime using the sysfs "same_filled_pages_enabled" attribute, e.g. +disabled at boot time by setting the ``same_filled_pages_enabled`` attribute +to 0, e.g. ``zswap.same_filled_pages_enabled=0``. It can also be enabled and +disabled at runtime using the sysfs ``same_filled_pages_enabled`` +attribute, e.g.:: -echo 1 > /sys/module/zswap/parameters/same_filled_pages_enabled + echo 1 > /sys/module/zswap/parameters/same_filled_pages_enabled When zswap same-filled page identification is disabled at runtime, it will stop checking for the same-value filled pages during store operation. However, the |