diff options
Diffstat (limited to 'Documentation')
29 files changed, 385 insertions, 592 deletions
diff --git a/Documentation/ABI/testing/sysfs-kernel-mm-damon b/Documentation/ABI/testing/sysfs-kernel-mm-damon index 2744f21b5a6b..334352d198f8 100644 --- a/Documentation/ABI/testing/sysfs-kernel-mm-damon +++ b/Documentation/ABI/testing/sysfs-kernel-mm-damon @@ -29,8 +29,10 @@ Description: Writing 'on' or 'off' to this file makes the kdamond starts or file updates contents of schemes stats files of the kdamond. Writing 'update_schemes_tried_regions' to the file updates contents of 'tried_regions' directory of every scheme directory - of this kdamond. Writing 'clear_schemes_tried_regions' to the - file removes contents of the 'tried_regions' directory. + of this kdamond. Writing 'update_schemes_tried_bytes' to the + file updates only '.../tried_regions/total_bytes' files of this + kdamond. Writing 'clear_schemes_tried_regions' to the file + removes contents of the 'tried_regions' directory. What: /sys/kernel/mm/damon/admin/kdamonds/<K>/pid Date: Mar 2022 @@ -269,8 +271,10 @@ What: /sys/kernel/mm/damon/admin/kdamonds/<K>/contexts/<C>/schemes/<S>/filters/ Date: Dec 2022 Contact: SeongJae Park <sj@kernel.org> Description: Writing to and reading from this file sets and gets the type of - the memory of the interest. 'anon' for anonymous pages, or - 'memcg' for specific memory cgroup can be written and read. + the memory of the interest. 'anon' for anonymous pages, + 'memcg' for specific memory cgroup, 'addr' for address range + (an open-ended interval), or 'target' for DAMON monitoring + target can be written and read. What: /sys/kernel/mm/damon/admin/kdamonds/<K>/contexts/<C>/schemes/<S>/filters/<F>/memcg_path Date: Dec 2022 @@ -279,6 +283,27 @@ Description: If 'memcg' is written to the 'type' file, writing to and reading from this file sets and gets the path to the memory cgroup of the interest. +What: /sys/kernel/mm/damon/admin/kdamonds/<K>/contexts/<C>/schemes/<S>/filters/<F>/addr_start +Date: Jul 2023 +Contact: SeongJae Park <sj@kernel.org> +Description: If 'addr' is written to the 'type' file, writing to or reading + from this file sets or gets the start address of the address + range for the filter. + +What: /sys/kernel/mm/damon/admin/kdamonds/<K>/contexts/<C>/schemes/<S>/filters/<F>/addr_end +Date: Jul 2023 +Contact: SeongJae Park <sj@kernel.org> +Description: If 'addr' is written to the 'type' file, writing to or reading + from this file sets or gets the end address of the address + range for the filter. + +What: /sys/kernel/mm/damon/admin/kdamonds/<K>/contexts/<C>/schemes/<S>/filters/<F>/target_idx +Date: Dec 2022 +Contact: SeongJae Park <sj@kernel.org> +Description: If 'target' is written to the 'type' file, writing to or + reading from this file sets or gets the index of the DAMON + monitoring target of the interest. + What: /sys/kernel/mm/damon/admin/kdamonds/<K>/contexts/<C>/schemes/<S>/filters/<F>/matching Date: Dec 2022 Contact: SeongJae Park <sj@kernel.org> @@ -317,6 +342,13 @@ Contact: SeongJae Park <sj@kernel.org> Description: Reading this file returns the number of the exceed events of the scheme's quotas. +What: /sys/kernel/mm/damon/admin/kdamonds/<K>/contexts/<C>/schemes/<S>/tried_regions/total_bytes +Date: Jul 2023 +Contact: SeongJae Park <sj@kernel.org> +Description: Reading this file returns the total amount of memory that + corresponding DAMON-based Operation Scheme's action has tried + to be applied. + What: /sys/kernel/mm/damon/admin/kdamonds/<K>/contexts/<C>/schemes/<S>/tried_regions/<R>/start Date: Oct 2022 Contact: SeongJae Park <sj@kernel.org> diff --git a/Documentation/ABI/testing/sysfs-memory-page-offline b/Documentation/ABI/testing/sysfs-memory-page-offline index e14703f12fdf..00f4e35f916f 100644 --- a/Documentation/ABI/testing/sysfs-memory-page-offline +++ b/Documentation/ABI/testing/sysfs-memory-page-offline @@ -10,7 +10,7 @@ Description: dropping it if possible. The kernel will then be placed on the bad page list and never be reused. - The offlining is done in kernel specific granuality. + The offlining is done in kernel specific granularity. Normally it's the base page size of the kernel, but this might change. @@ -35,7 +35,7 @@ Description: to access this page assuming it's poisoned by the hardware. - The offlining is done in kernel specific granuality. + The offlining is done in kernel specific granularity. Normally it's the base page size of the kernel, but this might change. diff --git a/Documentation/admin-guide/cgroup-v1/memory.rst b/Documentation/admin-guide/cgroup-v1/memory.rst index fabaad3fd9c2..8d3afeede10e 100644 --- a/Documentation/admin-guide/cgroup-v1/memory.rst +++ b/Documentation/admin-guide/cgroup-v1/memory.rst @@ -92,8 +92,6 @@ Brief summary of control files. memory.oom_control set/show oom controls. memory.numa_stat show the number of memory usage per numa node - memory.kmem.limit_in_bytes This knob is deprecated and writing to - it will return -ENOTSUPP. memory.kmem.usage_in_bytes show current kernel memory allocation memory.kmem.failcnt show the number of kernel memory usage hits limits diff --git a/Documentation/admin-guide/kdump/vmcoreinfo.rst b/Documentation/admin-guide/kdump/vmcoreinfo.rst index f8ebb63b6c5d..599e8d3bcbc3 100644 --- a/Documentation/admin-guide/kdump/vmcoreinfo.rst +++ b/Documentation/admin-guide/kdump/vmcoreinfo.rst @@ -141,8 +141,8 @@ nodemask_t The size of a nodemask_t type. Used to compute the number of online nodes. -(page, flags|_refcount|mapping|lru|_mapcount|private|compound_dtor|compound_order|compound_head) -------------------------------------------------------------------------------------------------- +(page, flags|_refcount|mapping|lru|_mapcount|private|compound_order|compound_head) +---------------------------------------------------------------------------------- User-space tools compute their values based on the offset of these variables. The variables are used when excluding unnecessary pages. @@ -325,8 +325,8 @@ NR_FREE_PAGES On linux-2.6.21 or later, the number of free pages is in vm_stat[NR_FREE_PAGES]. Used to get the number of free pages. -PG_lru|PG_private|PG_swapcache|PG_swapbacked|PG_slab|PG_hwpoision|PG_head_mask ------------------------------------------------------------------------------- +PG_lru|PG_private|PG_swapcache|PG_swapbacked|PG_slab|PG_hwpoision|PG_head_mask|PG_hugetlb +----------------------------------------------------------------------------------------- Page attributes. These flags are used to filter various unnecessary for dumping pages. @@ -338,12 +338,6 @@ More page attributes. These flags are used to filter various unnecessary for dumping pages. -HUGETLB_PAGE_DTOR ------------------ - -The HUGETLB_PAGE_DTOR flag denotes hugetlbfs pages. Makedumpfile -excludes these pages. - x86_64 ====== diff --git a/Documentation/admin-guide/mm/damon/usage.rst b/Documentation/admin-guide/mm/damon/usage.rst index 2d495fa85a0e..084f0a32b421 100644 --- a/Documentation/admin-guide/mm/damon/usage.rst +++ b/Documentation/admin-guide/mm/damon/usage.rst @@ -87,7 +87,7 @@ comma (","). :: │ │ │ │ │ │ │ filters/nr_filters │ │ │ │ │ │ │ │ 0/type,matching,memcg_id │ │ │ │ │ │ │ stats/nr_tried,sz_tried,nr_applied,sz_applied,qt_exceeds - │ │ │ │ │ │ │ tried_regions/ + │ │ │ │ │ │ │ tried_regions/total_bytes │ │ │ │ │ │ │ │ 0/start,end,nr_accesses,age │ │ │ │ │ │ │ │ ... │ │ │ │ │ │ ... @@ -127,14 +127,18 @@ in the state. Writing ``commit`` to the ``state`` file makes kdamond reads the user inputs in the sysfs files except ``state`` file again. Writing ``update_schemes_stats`` to ``state`` file updates the contents of stats files for each DAMON-based operation scheme of the kdamond. For details of the -stats, please refer to :ref:`stats section <sysfs_schemes_stats>`. Writing -``update_schemes_tried_regions`` to ``state`` file updates the DAMON-based -operation scheme action tried regions directory for each DAMON-based operation -scheme of the kdamond. Writing ``clear_schemes_tried_regions`` to ``state`` -file clears the DAMON-based operating scheme action tried regions directory for -each DAMON-based operation scheme of the kdamond. For details of the -DAMON-based operation scheme action tried regions directory, please refer to -:ref:`tried_regions section <sysfs_schemes_tried_regions>`. +stats, please refer to :ref:`stats section <sysfs_schemes_stats>`. + +Writing ``update_schemes_tried_regions`` to ``state`` file updates the +DAMON-based operation scheme action tried regions directory for each +DAMON-based operation scheme of the kdamond. Writing +``update_schemes_tried_bytes`` to ``state`` file updates only +``.../tried_regions/total_bytes`` files. Writing +``clear_schemes_tried_regions`` to ``state`` file clears the DAMON-based +operating scheme action tried regions directory for each DAMON-based operation +scheme of the kdamond. For details of the DAMON-based operation scheme action +tried regions directory, please refer to :ref:`tried_regions section +<sysfs_schemes_tried_regions>`. If the state is ``on``, reading ``pid`` shows the pid of the kdamond thread. @@ -359,15 +363,21 @@ number (``N``) to the file creates the number of child directories named ``0`` to ``N-1``. Each directory represents each filter. The filters are evaluated in the numeric order. -Each filter directory contains three files, namely ``type``, ``matcing``, and -``memcg_path``. You can write one of two special keywords, ``anon`` for -anonymous pages, or ``memcg`` for specific memory cgroup filtering. In case of -the memory cgroup filtering, you can specify the memory cgroup of the interest -by writing the path of the memory cgroup from the cgroups mount point to -``memcg_path`` file. You can write ``Y`` or ``N`` to ``matching`` file to -filter out pages that does or does not match to the type, respectively. Then, -the scheme's action will not be applied to the pages that specified to be -filtered out. +Each filter directory contains six files, namely ``type``, ``matcing``, +``memcg_path``, ``addr_start``, ``addr_end``, and ``target_idx``. To ``type`` +file, you can write one of four special keywords: ``anon`` for anonymous pages, +``memcg`` for specific memory cgroup, ``addr`` for specific address range (an +open-ended interval), or ``target`` for specific DAMON monitoring target +filtering. In case of the memory cgroup filtering, you can specify the memory +cgroup of the interest by writing the path of the memory cgroup from the +cgroups mount point to ``memcg_path`` file. In case of the address range +filtering, you can specify the start and end address of the range to +``addr_start`` and ``addr_end`` files, respectively. For the DAMON monitoring +target filtering, you can specify the index of the target between the list of +the DAMON context's monitoring targets list to ``target_idx`` file. You can +write ``Y`` or ``N`` to ``matching`` file to filter out pages that does or does +not match to the type, respectively. Then, the scheme's action will not be +applied to the pages that specified to be filtered out. For example, below restricts a DAMOS action to be applied to only non-anonymous pages of all memory cgroups except ``/having_care_already``.:: @@ -381,8 +391,14 @@ pages of all memory cgroups except ``/having_care_already``.:: echo /having_care_already > 1/memcg_path echo N > 1/matching -Note that filters are currently supported only when ``paddr`` -`implementation <sysfs_contexts>` is being used. +Note that ``anon`` and ``memcg`` filters are currently supported only when +``paddr`` `implementation <sysfs_contexts>` is being used. + +Also, memory regions that are filtered out by ``addr`` or ``target`` filters +are not counted as the scheme has tried to those, while regions that filtered +out by other type filters are counted as the scheme has tried to. The +difference is applied to :ref:`stats <damos_stats>` and +:ref:`tried regions <sysfs_schemes_tried_regions>`. .. _sysfs_schemes_stats: @@ -406,13 +422,21 @@ stats by writing a special keyword, ``update_schemes_stats`` to the relevant schemes/<N>/tried_regions/ -------------------------- +This directory initially has one file, ``total_bytes``. + When a special keyword, ``update_schemes_tried_regions``, is written to the -relevant ``kdamonds/<N>/state`` file, DAMON creates directories named integer -starting from ``0`` under this directory. Each directory contains files -exposing detailed information about each of the memory region that the -corresponding scheme's ``action`` has tried to be applied under this directory, -during next :ref:`aggregation interval <sysfs_monitoring_attrs>`. The -information includes address range, ``nr_accesses``, and ``age`` of the region. +relevant ``kdamonds/<N>/state`` file, DAMON updates the ``total_bytes`` file so +that reading it returns the total size of the scheme tried regions, and creates +directories named integer starting from ``0`` under this directory. Each +directory contains files exposing detailed information about each of the memory +region that the corresponding scheme's ``action`` has tried to be applied under +this directory, during next :ref:`aggregation interval +<sysfs_monitoring_attrs>`. The information includes address range, +``nr_accesses``, and ``age`` of the region. + +Writing ``update_schemes_tried_bytes`` to the relevant ``kdamonds/<N>/state`` +file will only update the ``total_bytes`` file, and will not create the +subdirectories. The directories will be removed when another special keyword, ``clear_schemes_tried_regions``, is written to the relevant diff --git a/Documentation/admin-guide/mm/ksm.rst b/Documentation/admin-guide/mm/ksm.rst index 7626392fe82c..776f244bdae4 100644 --- a/Documentation/admin-guide/mm/ksm.rst +++ b/Documentation/admin-guide/mm/ksm.rst @@ -159,6 +159,8 @@ The effectiveness of KSM and MADV_MERGEABLE is shown in ``/sys/kernel/mm/ksm/``: general_profit how effective is KSM. The calculation is explained below. +pages_scanned + how many pages are being scanned for ksm pages_shared how many shared pages are being used pages_sharing @@ -173,6 +175,13 @@ stable_node_chains the number of KSM pages that hit the ``max_page_sharing`` limit stable_node_dups number of duplicated KSM pages +ksm_zero_pages + how many zero pages that are still mapped into processes were mapped by + KSM when deduplicating. + +When ``use_zero_pages`` is/was enabled, the sum of ``pages_sharing`` + +``ksm_zero_pages`` represents the actual number of pages saved by KSM. +if ``use_zero_pages`` has never been enabled, ``ksm_zero_pages`` is 0. A high ratio of ``pages_sharing`` to ``pages_shared`` indicates good sharing, but a high ratio of ``pages_unshared`` to ``pages_sharing`` @@ -196,21 +205,25 @@ several times, which are unprofitable memory consumed. 1) How to determine whether KSM save memory or consume memory in system-wide range? Here is a simple approximate calculation for reference:: - general_profit =~ pages_sharing * sizeof(page) - (all_rmap_items) * + general_profit =~ ksm_saved_pages * sizeof(page) - (all_rmap_items) * sizeof(rmap_item); - where all_rmap_items can be easily obtained by summing ``pages_sharing``, - ``pages_shared``, ``pages_unshared`` and ``pages_volatile``. + where ksm_saved_pages equals to the sum of ``pages_sharing`` + + ``ksm_zero_pages`` of the system, and all_rmap_items can be easily + obtained by summing ``pages_sharing``, ``pages_shared``, ``pages_unshared`` + and ``pages_volatile``. 2) The KSM profit inner a single process can be similarly obtained by the following approximate calculation:: - process_profit =~ ksm_merging_pages * sizeof(page) - + process_profit =~ ksm_saved_pages * sizeof(page) - ksm_rmap_items * sizeof(rmap_item). - where ksm_merging_pages is shown under the directory ``/proc/<pid>/``, - and ksm_rmap_items is shown in ``/proc/<pid>/ksm_stat``. The process profit - is also shown in ``/proc/<pid>/ksm_stat`` as ksm_process_profit. + where ksm_saved_pages equals to the sum of ``ksm_merging_pages`` and + ``ksm_zero_pages``, both of which are shown under the directory + ``/proc/<pid>/ksm_stat``, and ksm_rmap_items is also shown in + ``/proc/<pid>/ksm_stat``. The process profit is also shown in + ``/proc/<pid>/ksm_stat`` as ksm_process_profit. From the perspective of application, a high ratio of ``ksm_rmap_items`` to ``ksm_merging_pages`` means a bad madvise-applied policy, so developers or diff --git a/Documentation/admin-guide/mm/memory-hotplug.rst b/Documentation/admin-guide/mm/memory-hotplug.rst index 1b02fe5807cc..2994958c7ce8 100644 --- a/Documentation/admin-guide/mm/memory-hotplug.rst +++ b/Documentation/admin-guide/mm/memory-hotplug.rst @@ -433,6 +433,18 @@ The following module parameters are currently defined: memory in a way that huge pages in bigger granularity cannot be formed on hotplugged memory. + + With value "force" it could result in memory + wastage due to memmap size limitations. For + example, if the memmap for a memory block + requires 1 MiB, but the pageblock size is 2 + MiB, 1 MiB of hotplugged memory will be wasted. + Note that there are still cases where the + feature cannot be enforced: for example, if the + memmap is smaller than a single page, or if the + architecture does not support the forced mode + in all configurations. + ``online_policy`` read-write: Set the basic policy used for automatic zone selection when onlining memory blocks without specifying a target zone. @@ -669,7 +681,7 @@ when still encountering permanently unmovable pages within ZONE_MOVABLE (-> BUG), memory offlining will keep retrying until it eventually succeeds. When offlining is triggered from user space, the offlining context can be -terminated by sending a fatal signal. A timeout based offlining can easily be +terminated by sending a signal. A timeout based offlining can easily be implemented via:: % timeout $TIMEOUT offline_block | failure_handling diff --git a/Documentation/admin-guide/mm/userfaultfd.rst b/Documentation/admin-guide/mm/userfaultfd.rst index 7c304e432205..4349a8c2b978 100644 --- a/Documentation/admin-guide/mm/userfaultfd.rst +++ b/Documentation/admin-guide/mm/userfaultfd.rst @@ -244,6 +244,21 @@ write-protected (so future writes will also result in a WP fault). These ioctls support a mode flag (``UFFDIO_COPY_MODE_WP`` or ``UFFDIO_CONTINUE_MODE_WP`` respectively) to configure the mapping this way. +Memory Poisioning Emulation +--------------------------- + +In response to a fault (either missing or minor), an action userspace can +take to "resolve" it is to issue a ``UFFDIO_POISON``. This will cause any +future faulters to either get a SIGBUS, or in KVM's case the guest will +receive an MCE as if there were hardware memory poisoning. + +This is used to emulate hardware memory poisoning. Imagine a VM running on a +machine which experiences a real hardware memory error. Later, we live migrate +the VM to another physical machine. Since we want the migration to be +transparent to the guest, we want that same address range to act as if it was +still poisoned, even though it's on a new physical host which ostensibly +doesn't have a memory error in the exact same spot. + QEMU/KVM ======== diff --git a/Documentation/admin-guide/mm/zswap.rst b/Documentation/admin-guide/mm/zswap.rst index c5c2c7dbb155..45b98390e938 100644 --- a/Documentation/admin-guide/mm/zswap.rst +++ b/Documentation/admin-guide/mm/zswap.rst @@ -49,7 +49,7 @@ compressed pool. Design ====== -Zswap receives pages for compression through the Frontswap API and is able to +Zswap receives pages for compression from the swap subsystem and is able to evict pages from its own compressed pool on an LRU basis and write them back to the backing swap device in the case that the compressed pool is full. @@ -70,19 +70,19 @@ means the compression ratio will always be 2:1 or worse (because of half-full zbud pages). The zsmalloc type zpool has a more complex compressed page storage method, and it can achieve greater storage densities. -When a swap page is passed from frontswap to zswap, zswap maintains a mapping +When a swap page is passed from swapout to zswap, zswap maintains a mapping of the swap entry, a combination of the swap type and swap offset, to the zpool handle that references that compressed swap page. This mapping is achieved with a red-black tree per swap type. The swap offset is the search key for the tree nodes. -During a page fault on a PTE that is a swap entry, frontswap calls the zswap -load function to decompress the page into the page allocated by the page fault -handler. +During a page fault on a PTE that is a swap entry, the swapin code calls the +zswap load function to decompress the page into the page allocated by the page +fault handler. Once there are no PTEs referencing a swap page stored in zswap (i.e. the count -in the swap_map goes to 0) the swap code calls the zswap invalidate function, -via frontswap, to free the compressed entry. +in the swap_map goes to 0) the swap code calls the zswap invalidate function +to free the compressed entry. Zswap seeks to be simple in its policies. Sysfs attributes allow for one user controlled policy: diff --git a/Documentation/block/biovecs.rst b/Documentation/block/biovecs.rst index ddb867e0185b..b9dc0c9dbee4 100644 --- a/Documentation/block/biovecs.rst +++ b/Documentation/block/biovecs.rst @@ -134,6 +134,7 @@ Usage of helpers: bio_for_each_bvec_all() bio_first_bvec_all() bio_first_page_all() + bio_first_folio_all() bio_last_bvec_all() * The following helpers iterate over single-page segment. The passed 'struct diff --git a/Documentation/core-api/cachetlb.rst b/Documentation/core-api/cachetlb.rst index 5c0552e78c58..889fc84ccd1b 100644 --- a/Documentation/core-api/cachetlb.rst +++ b/Documentation/core-api/cachetlb.rst @@ -88,13 +88,17 @@ changes occur: This is used primarily during fault processing. -5) ``void update_mmu_cache(struct vm_area_struct *vma, - unsigned long address, pte_t *ptep)`` +5) ``void update_mmu_cache_range(struct vm_fault *vmf, + struct vm_area_struct *vma, unsigned long address, pte_t *ptep, + unsigned int nr)`` - At the end of every page fault, this routine is invoked to - tell the architecture specific code that a translation - now exists at virtual address "address" for address space - "vma->vm_mm", in the software page tables. + At the end of every page fault, this routine is invoked to tell + the architecture specific code that translations now exists + in the software page tables for address space "vma->vm_mm" + at virtual address "address" for "nr" consecutive pages. + + This routine is also invoked in various other places which pass + a NULL "vmf". A port may use this information in any way it so chooses. For example, it could use this event to pre-load TLB @@ -269,7 +273,7 @@ maps this page at its virtual address. If D-cache aliasing is not an issue, these two routines may simply call memcpy/memset directly and do nothing more. - ``void flush_dcache_page(struct page *page)`` + ``void flush_dcache_folio(struct folio *folio)`` This routines must be called when: @@ -277,7 +281,7 @@ maps this page at its virtual address. and / or in high memory b) the kernel is about to read from a page cache page and user space shared/writable mappings of this page potentially exist. Note - that {get,pin}_user_pages{_fast} already call flush_dcache_page + that {get,pin}_user_pages{_fast} already call flush_dcache_folio on any page found in the user address space and thus driver code rarely needs to take this into account. @@ -291,7 +295,7 @@ maps this page at its virtual address. The phrase "kernel writes to a page cache page" means, specifically, that the kernel executes store instructions that dirty data in that - page at the page->virtual mapping of that page. It is important to + page at the kernel virtual mapping of that page. It is important to flush here to handle D-cache aliasing, to make sure these kernel stores are visible to user space mappings of that page. @@ -302,21 +306,22 @@ maps this page at its virtual address. If D-cache aliasing is not an issue, this routine may simply be defined as a nop on that architecture. - There is a bit set aside in page->flags (PG_arch_1) as "architecture + There is a bit set aside in folio->flags (PG_arch_1) as "architecture private". The kernel guarantees that, for pagecache pages, it will clear this bit when such a page first enters the pagecache. - This allows these interfaces to be implemented much more efficiently. - It allows one to "defer" (perhaps indefinitely) the actual flush if - there are currently no user processes mapping this page. See sparc64's - flush_dcache_page and update_mmu_cache implementations for an example - of how to go about doing this. + This allows these interfaces to be implemented much more + efficiently. It allows one to "defer" (perhaps indefinitely) the + actual flush if there are currently no user processes mapping this + page. See sparc64's flush_dcache_folio and update_mmu_cache_range + implementations for an example of how to go about doing this. - The idea is, first at flush_dcache_page() time, if page_file_mapping() - returns a mapping, and mapping_mapped on that mapping returns %false, - just mark the architecture private page flag bit. Later, in - update_mmu_cache(), a check is made of this flag bit, and if set the - flush is done and the flag bit is cleared. + The idea is, first at flush_dcache_folio() time, if + folio_flush_mapping() returns a mapping, and mapping_mapped() on that + mapping returns %false, just mark the architecture private page + flag bit. Later, in update_mmu_cache_range(), a check is made + of this flag bit, and if set the flush is done and the flag bit + is cleared. .. important:: @@ -326,12 +331,6 @@ maps this page at its virtual address. dirty. Again, see sparc64 for examples of how to deal with this. - ``void flush_dcache_folio(struct folio *folio)`` - This function is called under the same circumstances as - flush_dcache_page(). It allows the architecture to - optimise for flushing the entire folio of pages instead - of flushing one page at a time. - ``void copy_to_user_page(struct vm_area_struct *vma, struct page *page, unsigned long user_vaddr, void *dst, void *src, int len)`` ``void copy_from_user_page(struct vm_area_struct *vma, struct page *page, @@ -352,7 +351,7 @@ maps this page at its virtual address. When the kernel needs to access the contents of an anonymous page, it calls this function (currently only - get_user_pages()). Note: flush_dcache_page() deliberately + get_user_pages()). Note: flush_dcache_folio() deliberately doesn't work for an anonymous page. The default implementation is a nop (and should remain so for all coherent architectures). For incoherent architectures, it should flush @@ -369,7 +368,7 @@ maps this page at its virtual address. ``void flush_icache_page(struct vm_area_struct *vma, struct page *page)`` All the functionality of flush_icache_page can be implemented in - flush_dcache_page and update_mmu_cache. In the future, the hope + flush_dcache_folio and update_mmu_cache_range. In the future, the hope is to remove this interface completely. The final category of APIs is for I/O to deliberately aliased address diff --git a/Documentation/core-api/mm-api.rst b/Documentation/core-api/mm-api.rst index f5dde5bceaea..2d091c873d1e 100644 --- a/Documentation/core-api/mm-api.rst +++ b/Documentation/core-api/mm-api.rst @@ -115,3 +115,28 @@ More Memory Management Functions .. kernel-doc:: include/linux/mmzone.h .. kernel-doc:: mm/util.c :functions: folio_mapping + +.. kernel-doc:: mm/rmap.c +.. kernel-doc:: mm/migrate.c +.. kernel-doc:: mm/mmap.c +.. kernel-doc:: mm/kmemleak.c +.. #kernel-doc:: mm/hmm.c (build warnings) +.. kernel-doc:: mm/memremap.c +.. kernel-doc:: mm/hugetlb.c +.. kernel-doc:: mm/swap.c +.. kernel-doc:: mm/zpool.c +.. kernel-doc:: mm/memcontrol.c +.. #kernel-doc:: mm/memory-tiers.c (build warnings) +.. kernel-doc:: mm/shmem.c +.. kernel-doc:: mm/migrate_device.c +.. #kernel-doc:: mm/nommu.c (duplicates kernel-doc from other files) +.. kernel-doc:: mm/mapping_dirty_helpers.c +.. #kernel-doc:: mm/memory-failure.c (build warnings) +.. kernel-doc:: mm/percpu.c +.. kernel-doc:: mm/maccess.c +.. kernel-doc:: mm/vmscan.c +.. kernel-doc:: mm/memory_hotplug.c +.. kernel-doc:: mm/mmu_notifier.c +.. kernel-doc:: mm/balloon_compaction.c +.. kernel-doc:: mm/huge_memory.c +.. kernel-doc:: mm/io-mapping.c diff --git a/Documentation/features/vm/TLB/arch-support.txt b/Documentation/features/vm/TLB/arch-support.txt index 7f049c251a79..76208db88f3b 100644 --- a/Documentation/features/vm/TLB/arch-support.txt +++ b/Documentation/features/vm/TLB/arch-support.txt @@ -9,7 +9,7 @@ | alpha: | TODO | | arc: | TODO | | arm: | TODO | - | arm64: | N/A | + | arm64: | ok | | csky: | TODO | | hexagon: | TODO | | ia64: | TODO | diff --git a/Documentation/filesystems/locking.rst b/Documentation/filesystems/locking.rst index b7e5a3841aa4..2fd01b9aaced 100644 --- a/Documentation/filesystems/locking.rst +++ b/Documentation/filesystems/locking.rst @@ -636,26 +636,29 @@ vm_operations_struct prototypes:: - void (*open)(struct vm_area_struct*); - void (*close)(struct vm_area_struct*); - vm_fault_t (*fault)(struct vm_area_struct*, struct vm_fault *); + void (*open)(struct vm_area_struct *); + void (*close)(struct vm_area_struct *); + vm_fault_t (*fault)(struct vm_fault *); + vm_fault_t (*huge_fault)(struct vm_fault *, unsigned int order); + vm_fault_t (*map_pages)(struct vm_fault *, pgoff_t start, pgoff_t end); vm_fault_t (*page_mkwrite)(struct vm_area_struct *, struct vm_fault *); vm_fault_t (*pfn_mkwrite)(struct vm_area_struct *, struct vm_fault *); int (*access)(struct vm_area_struct *, unsigned long, void*, int, int); locking rules: -============= ========= =========================== +============= ========== =========================== ops mmap_lock PageLocked(page) -============= ========= =========================== -open: yes -close: yes -fault: yes can return with page locked -map_pages: read -page_mkwrite: yes can return with page locked -pfn_mkwrite: yes -access: yes -============= ========= =========================== +============= ========== =========================== +open: write +close: read/write +fault: read can return with page locked +huge_fault: maybe-read +map_pages: maybe-read +page_mkwrite: read can return with page locked +pfn_mkwrite: read +access: read +============= ========== =========================== ->fault() is called when a previously not present pte is about to be faulted in. The filesystem must find and return the page associated with the passed in @@ -665,11 +668,18 @@ then ensure the page is not already truncated (invalidate_lock will block subsequent truncate), and then return with VM_FAULT_LOCKED, and the page locked. The VM will unlock the page. +->huge_fault() is called when there is no PUD or PMD entry present. This +gives the filesystem the opportunity to install a PUD or PMD sized page. +Filesystems can also use the ->fault method to return a PMD sized page, +so implementing this function may not be necessary. In particular, +filesystems should not call filemap_fault() from ->huge_fault(). +The mmap_lock may not be held when this method is called. + ->map_pages() is called when VM asks to map easy accessible pages. Filesystem should find and map pages associated with offsets from "start_pgoff" till "end_pgoff". ->map_pages() is called with the RCU lock held and must not block. If it's not possible to reach a page without blocking, -filesystem should skip it. Filesystem should use do_set_pte() to setup +filesystem should skip it. Filesystem should use set_pte_range() to setup page table entry. Pointer to entry associated with the page is passed in "pte" field in vm_fault structure. Pointers to entries for other offsets should be calculated relative to "pte". diff --git a/Documentation/filesystems/porting.rst b/Documentation/filesystems/porting.rst index 0f5da78ef4f9..98969d713e2e 100644 --- a/Documentation/filesystems/porting.rst +++ b/Documentation/filesystems/porting.rst @@ -938,3 +938,14 @@ file pointer instead of struct dentry pointer. d_tmpfile() is similarly changed to simplify callers. The passed file is in a non-open state and on success must be opened before returning (e.g. by calling finish_open_simple()). + +--- + +**mandatory** + +Calling convention for ->huge_fault has changed. It now takes a page +order instead of an enum page_entry_size, and it may be called without the +mmap_lock held. All in-tree users have been audited and do not seem to +depend on the mmap_lock being held, but out of tree users should verify +for themselves. If they do need it, they can return VM_FAULT_RETRY to +be called with the mmap_lock held. diff --git a/Documentation/mm/damon/design.rst b/Documentation/mm/damon/design.rst index 4bfdf1d30c4a..a20383d01a95 100644 --- a/Documentation/mm/damon/design.rst +++ b/Documentation/mm/damon/design.rst @@ -380,12 +380,24 @@ number of filters for each scheme. Each filter specifies the type of target memory, and whether it should exclude the memory of the type (filter-out), or all except the memory of the type (filter-in). -As of this writing, anonymous page type and memory cgroup type are supported by -the feature. Some filter target types can require additional arguments. For -example, the memory cgroup filter type asks users to specify the file path of -the memory cgroup for the filter. Hence, users can apply specific schemes to -only anonymous pages, non-anonymous pages, pages of specific cgroups, all pages -excluding those of specific cgroups, and any combination of those. +Currently, anonymous page, memory cgroup, address range, and DAMON monitoring +target type filters are supported by the feature. Some filter target types +require additional arguments. The memory cgroup filter type asks users to +specify the file path of the memory cgroup for the filter. The address range +type asks the start and end addresses of the range. The DAMON monitoring +target type asks the index of the target from the context's monitoring targets +list. Hence, users can apply specific schemes to only anonymous pages, +non-anonymous pages, pages of specific cgroups, all pages excluding those of +specific cgroups, pages in specific address range, pages in specific DAMON +monitoring targets, and any combination of those. + +To handle filters efficiently, the address range and DAMON monitoring target +type filters are handled by the core layer, while others are handled by +operations set. If a memory region is filtered by a core layer-handled filter, +it is not counted as the scheme has tried to the region. In contrast, if a +memory regions is filtered by an operations set layer-handled filter, it is +counted as the scheme has tried. The difference in accounting leads to changes +in the statistics. Application Programming Interface diff --git a/Documentation/mm/frontswap.rst b/Documentation/mm/frontswap.rst deleted file mode 100644 index c892412988af..000000000000 --- a/Documentation/mm/frontswap.rst +++ /dev/null @@ -1,264 +0,0 @@ -========= -Frontswap -========= - -Frontswap provides a "transcendent memory" interface for swap pages. -In some environments, dramatic performance savings may be obtained because -swapped pages are saved in RAM (or a RAM-like device) instead of a swap disk. - -.. _Transcendent memory in a nutshell: https://lwn.net/Articles/454795/ - -Frontswap is so named because it can be thought of as the opposite of -a "backing" store for a swap device. The storage is assumed to be -a synchronous concurrency-safe page-oriented "pseudo-RAM device" conforming -to the requirements of transcendent memory (such as Xen's "tmem", or -in-kernel compressed memory, aka "zcache", or future RAM-like devices); -this pseudo-RAM device is not directly accessible or addressable by the -kernel and is of unknown and possibly time-varying size. The driver -links itself to frontswap by calling frontswap_register_ops to set the -frontswap_ops funcs appropriately and the functions it provides must -conform to certain policies as follows: - -An "init" prepares the device to receive frontswap pages associated -with the specified swap device number (aka "type"). A "store" will -copy the page to transcendent memory and associate it with the type and -offset associated with the page. A "load" will copy the page, if found, -from transcendent memory into kernel memory, but will NOT remove the page -from transcendent memory. An "invalidate_page" will remove the page -from transcendent memory and an "invalidate_area" will remove ALL pages -associated with the swap type (e.g., like swapoff) and notify the "device" -to refuse further stores with that swap type. - -Once a page is successfully stored, a matching load on the page will normally -succeed. So when the kernel finds itself in a situation where it needs -to swap out a page, it first attempts to use frontswap. If the store returns -success, the data has been successfully saved to transcendent memory and -a disk write and, if the data is later read back, a disk read are avoided. -If a store returns failure, transcendent memory has rejected the data, and the -page can be written to swap as usual. - -Note that if a page is stored and the page already exists in transcendent memory -(a "duplicate" store), either the store succeeds and the data is overwritten, -or the store fails AND the page is invalidated. This ensures stale data may -never be obtained from frontswap. - -If properly configured, monitoring of frontswap is done via debugfs in -the `/sys/kernel/debug/frontswap` directory. The effectiveness of -frontswap can be measured (across all swap devices) with: - -``failed_stores`` - how many store attempts have failed - -``loads`` - how many loads were attempted (all should succeed) - -``succ_stores`` - how many store attempts have succeeded - -``invalidates`` - how many invalidates were attempted - -A backend implementation may provide additional metrics. - -FAQ -=== - -* Where's the value? - -When a workload starts swapping, performance falls through the floor. -Frontswap significantly increases performance in many such workloads by -providing a clean, dynamic interface to read and write swap pages to -"transcendent memory" that is otherwise not directly addressable to the kernel. -This interface is ideal when data is transformed to a different form -and size (such as with compression) or secretly moved (as might be -useful for write-balancing for some RAM-like devices). Swap pages (and -evicted page-cache pages) are a great use for this kind of slower-than-RAM- -but-much-faster-than-disk "pseudo-RAM device". - -Frontswap with a fairly small impact on the kernel, -provides a huge amount of flexibility for more dynamic, flexible RAM -utilization in various system configurations: - -In the single kernel case, aka "zcache", pages are compressed and -stored in local memory, thus increasing the total anonymous pages -that can be safely kept in RAM. Zcache essentially trades off CPU -cycles used in compression/decompression for better memory utilization. -Benchmarks have shown little or no impact when memory pressure is -low while providing a significant performance improvement (25%+) -on some workloads under high memory pressure. - -"RAMster" builds on zcache by adding "peer-to-peer" transcendent memory -support for clustered systems. Frontswap pages are locally compressed -as in zcache, but then "remotified" to another system's RAM. This -allows RAM to be dynamically load-balanced back-and-forth as needed, -i.e. when system A is overcommitted, it can swap to system B, and -vice versa. RAMster can also be configured as a memory server so -many servers in a cluster can swap, dynamically as needed, to a single -server configured with a large amount of RAM... without pre-configuring -how much of the RAM is available for each of the clients! - -In the virtual case, the whole point of virtualization is to statistically -multiplex physical resources across the varying demands of multiple -virtual machines. This is really hard to do with RAM and efforts to do -it well with no kernel changes have essentially failed (except in some -well-publicized special-case workloads). -Specifically, the Xen Transcendent Memory backend allows otherwise -"fallow" hypervisor-owned RAM to not only be "time-shared" between multiple -virtual machines, but the pages can be compressed and deduplicated to -optimize RAM utilization. And when guest OS's are induced to surrender -underutilized RAM (e.g. with "selfballooning"), sudden unexpected -memory pressure may result in swapping; frontswap allows those pages -to be swapped to and from hypervisor RAM (if overall host system memory -conditions allow), thus mitigating the potentially awful performance impact -of unplanned swapping. - -A KVM implementation is underway and has been RFC'ed to lkml. And, -using frontswap, investigation is also underway on the use of NVM as -a memory extension technology. - -* Sure there may be performance advantages in some situations, but - what's the space/time overhead of frontswap? - -If CONFIG_FRONTSWAP is disabled, every frontswap hook compiles into -nothingness and the only overhead is a few extra bytes per swapon'ed -swap device. If CONFIG_FRONTSWAP is enabled but no frontswap "backend" -registers, there is one extra global variable compared to zero for -every swap page read or written. If CONFIG_FRONTSWAP is enabled -AND a frontswap backend registers AND the backend fails every "store" -request (i.e. provides no memory despite claiming it might), -CPU overhead is still negligible -- and since every frontswap fail -precedes a swap page write-to-disk, the system is highly likely -to be I/O bound and using a small fraction of a percent of a CPU -will be irrelevant anyway. - -As for space, if CONFIG_FRONTSWAP is enabled AND a frontswap backend -registers, one bit is allocated for every swap page for every swap -device that is swapon'd. This is added to the EIGHT bits (which -was sixteen until about 2.6.34) that the kernel already allocates -for every swap page for every swap device that is swapon'd. (Hugh -Dickins has observed that frontswap could probably steal one of -the existing eight bits, but let's worry about that minor optimization -later.) For very large swap disks (which are rare) on a standard -4K pagesize, this is 1MB per 32GB swap. - -When swap pages are stored in transcendent memory instead of written -out to disk, there is a side effect that this may create more memory -pressure that can potentially outweigh the other advantages. A -backend, such as zcache, must implement policies to carefully (but -dynamically) manage memory limits to ensure this doesn't happen. - -* OK, how about a quick overview of what this frontswap patch does - in terms that a kernel hacker can grok? - -Let's assume that a frontswap "backend" has registered during -kernel initialization; this registration indicates that this -frontswap backend has access to some "memory" that is not directly -accessible by the kernel. Exactly how much memory it provides is -entirely dynamic and random. - -Whenever a swap-device is swapon'd frontswap_init() is called, -passing the swap device number (aka "type") as a parameter. -This notifies frontswap to expect attempts to "store" swap pages -associated with that number. - -Whenever the swap subsystem is readying a page to write to a swap -device (c.f swap_writepage()), frontswap_store is called. Frontswap -consults with the frontswap backend and if the backend says it does NOT -have room, frontswap_store returns -1 and the kernel swaps the page -to the swap device as normal. Note that the response from the frontswap -backend is unpredictable to the kernel; it may choose to never accept a -page, it could accept every ninth page, or it might accept every -page. But if the backend does accept a page, the data from the page -has already been copied and associated with the type and offset, -and the backend guarantees the persistence of the data. In this case, -frontswap sets a bit in the "frontswap_map" for the swap device -corresponding to the page offset on the swap device to which it would -otherwise have written the data. - -When the swap subsystem needs to swap-in a page (swap_readpage()), -it first calls frontswap_load() which checks the frontswap_map to -see if the page was earlier accepted by the frontswap backend. If -it was, the page of data is filled from the frontswap backend and -the swap-in is complete. If not, the normal swap-in code is -executed to obtain the page of data from the real swap device. - -So every time the frontswap backend accepts a page, a swap device read -and (potentially) a swap device write are replaced by a "frontswap backend -store" and (possibly) a "frontswap backend loads", which are presumably much -faster. - -* Can't frontswap be configured as a "special" swap device that is - just higher priority than any real swap device (e.g. like zswap, - or maybe swap-over-nbd/NFS)? - -No. First, the existing swap subsystem doesn't allow for any kind of -swap hierarchy. Perhaps it could be rewritten to accommodate a hierarchy, -but this would require fairly drastic changes. Even if it were -rewritten, the existing swap subsystem uses the block I/O layer which -assumes a swap device is fixed size and any page in it is linearly -addressable. Frontswap barely touches the existing swap subsystem, -and works around the constraints of the block I/O subsystem to provide -a great deal of flexibility and dynamicity. - -For example, the acceptance of any swap page by the frontswap backend is -entirely unpredictable. This is critical to the definition of frontswap -backends because it grants completely dynamic discretion to the -backend. In zcache, one cannot know a priori how compressible a page is. -"Poorly" compressible pages can be rejected, and "poorly" can itself be -defined dynamically depending on current memory constraints. - -Further, frontswap is entirely synchronous whereas a real swap -device is, by definition, asynchronous and uses block I/O. The -block I/O layer is not only unnecessary, but may perform "optimizations" -that are inappropriate for a RAM-oriented device including delaying -the write of some pages for a significant amount of time. Synchrony is -required to ensure the dynamicity of the backend and to avoid thorny race -conditions that would unnecessarily and greatly complicate frontswap -and/or the block I/O subsystem. That said, only the initial "store" -and "load" operations need be synchronous. A separate asynchronous thread -is free to manipulate the pages stored by frontswap. For example, -the "remotification" thread in RAMster uses standard asynchronous -kernel sockets to move compressed frontswap pages to a remote machine. -Similarly, a KVM guest-side implementation could do in-guest compression -and use "batched" hypercalls. - -In a virtualized environment, the dynamicity allows the hypervisor -(or host OS) to do "intelligent overcommit". For example, it can -choose to accept pages only until host-swapping might be imminent, -then force guests to do their own swapping. - -There is a downside to the transcendent memory specifications for -frontswap: Since any "store" might fail, there must always be a real -slot on a real swap device to swap the page. Thus frontswap must be -implemented as a "shadow" to every swapon'd device with the potential -capability of holding every page that the swap device might have held -and the possibility that it might hold no pages at all. This means -that frontswap cannot contain more pages than the total of swapon'd -swap devices. For example, if NO swap device is configured on some -installation, frontswap is useless. Swapless portable devices -can still use frontswap but a backend for such devices must configure -some kind of "ghost" swap device and ensure that it is never used. - -* Why this weird definition about "duplicate stores"? If a page - has been previously successfully stored, can't it always be - successfully overwritten? - -Nearly always it can, but no, sometimes it cannot. Consider an example -where data is compressed and the original 4K page has been compressed -to 1K. Now an attempt is made to overwrite the page with data that -is non-compressible and so would take the entire 4K. But the backend -has no more space. In this case, the store must be rejected. Whenever -frontswap rejects a store that would overwrite, it also must invalidate -the old data and ensure that it is no longer accessible. Since the -swap subsystem then writes the new data to the read swap device, -this is the correct course of action to ensure coherency. - -* Why does the frontswap patch create the new include file swapfile.h? - -The frontswap code depends on some swap-subsystem-internal data -structures that have, over the years, moved back and forth between -static and global. This seemed a reasonable compromise: Define -them as global but declare them in a new include file that isn't -included by the large number of source files that include swap.h. - -Dan Magenheimer, last updated April 9, 2012 diff --git a/Documentation/mm/highmem.rst b/Documentation/mm/highmem.rst index c964e0848702..aefb03eb386e 100644 --- a/Documentation/mm/highmem.rst +++ b/Documentation/mm/highmem.rst @@ -206,4 +206,5 @@ Functions ========= .. kernel-doc:: include/linux/highmem.h +.. kernel-doc:: mm/highmem.c .. kernel-doc:: include/linux/highmem-internal.h diff --git a/Documentation/mm/hugetlbfs_reserv.rst b/Documentation/mm/hugetlbfs_reserv.rst index d9c2b0f01dcd..4914fbf07966 100644 --- a/Documentation/mm/hugetlbfs_reserv.rst +++ b/Documentation/mm/hugetlbfs_reserv.rst @@ -271,12 +271,12 @@ to the global reservation count (resv_huge_pages). Freeing Huge Pages ================== -Huge page freeing is performed by the routine free_huge_page(). This routine -is the destructor for hugetlbfs compound pages. As a result, it is only -passed a pointer to the page struct. When a huge page is freed, reservation -accounting may need to be performed. This would be the case if the page was -associated with a subpool that contained reserves, or the page is being freed -on an error path where a global reserve count must be restored. +Huge pages are freed by free_huge_folio(). It is only passed a pointer +to the folio as it is called from the generic MM code. When a huge page +is freed, reservation accounting may need to be performed. This would +be the case if the page was associated with a subpool that contained +reserves, or the page is being freed on an error path where a global +reserve count must be restored. The page->private field points to any subpool associated with the page. If the PagePrivate flag is set, it indicates the global reserve count should @@ -525,7 +525,7 @@ However, there are several instances where errors are encountered after a huge page is allocated but before it is instantiated. In this case, the page allocation has consumed the reservation and made the appropriate subpool, reservation map and global count adjustments. If the page is freed at this -time (before instantiation and clearing of PagePrivate), then free_huge_page +time (before instantiation and clearing of PagePrivate), then free_huge_folio will increment the global reservation count. However, the reservation map indicates the reservation was consumed. This resulting inconsistent state will cause the 'leak' of a reserved huge page. The global reserve count will diff --git a/Documentation/mm/index.rst b/Documentation/mm/index.rst index 5a94a921ea40..31d2ac306438 100644 --- a/Documentation/mm/index.rst +++ b/Documentation/mm/index.rst @@ -44,7 +44,6 @@ above structured documentation, or deleted if it has served its purpose. balance damon/index free_page_reporting - frontswap hmm hwpoison hugetlbfs_reserv diff --git a/Documentation/mm/split_page_table_lock.rst b/Documentation/mm/split_page_table_lock.rst index a834fad9de12..e4f6972eb6c0 100644 --- a/Documentation/mm/split_page_table_lock.rst +++ b/Documentation/mm/split_page_table_lock.rst @@ -58,7 +58,7 @@ Support of split page table lock by an architecture =================================================== There's no need in special enabling of PTE split page table lock: everything -required is done by pgtable_pte_page_ctor() and pgtable_pte_page_dtor(), which +required is done by pagetable_pte_ctor() and pagetable_pte_dtor(), which must be called on PTE table allocation / freeing. Make sure the architecture doesn't use slab allocator for page table @@ -68,8 +68,8 @@ This field shares storage with page->ptl. PMD split lock only makes sense if you have more than two page table levels. -PMD split lock enabling requires pgtable_pmd_page_ctor() call on PMD table -allocation and pgtable_pmd_page_dtor() on freeing. +PMD split lock enabling requires pagetable_pmd_ctor() call on PMD table +allocation and pagetable_pmd_dtor() on freeing. Allocation usually happens in pmd_alloc_one(), freeing in pmd_free() and pmd_free_tlb(), but make sure you cover all PMD table allocation / freeing @@ -77,7 +77,7 @@ paths: i.e X86_PAE preallocate few PMDs on pgd_alloc(). With everything in place you can set CONFIG_ARCH_ENABLE_SPLIT_PMD_PTLOCK. -NOTE: pgtable_pte_page_ctor() and pgtable_pmd_page_ctor() can fail -- it must +NOTE: pagetable_pte_ctor() and pagetable_pmd_ctor() can fail -- it must be handled properly. page->ptl @@ -97,7 +97,7 @@ trick: split lock with enabled DEBUG_SPINLOCK or DEBUG_LOCK_ALLOC, but costs one more cache line for indirect access; -The spinlock_t allocated in pgtable_pte_page_ctor() for PTE table and in -pgtable_pmd_page_ctor() for PMD table. +The spinlock_t allocated in pagetable_pte_ctor() for PTE table and in +pagetable_pmd_ctor() for PMD table. Please, never access page->ptl directly -- use appropriate helper. diff --git a/Documentation/mm/vmemmap_dedup.rst b/Documentation/mm/vmemmap_dedup.rst index a4b12ff906c4..c573e08b5043 100644 --- a/Documentation/mm/vmemmap_dedup.rst +++ b/Documentation/mm/vmemmap_dedup.rst @@ -210,6 +210,7 @@ the device (altmap). The following page sizes are supported in DAX: PAGE_SIZE (4K on x86_64), PMD_SIZE (2M on x86_64) and PUD_SIZE (1G on x86_64). +For powerpc equivalent details see Documentation/powerpc/vmemmap_dedup.rst The differences with HugeTLB are relatively minor. diff --git a/Documentation/mm/zsmalloc.rst b/Documentation/mm/zsmalloc.rst index a3c26d587752..76902835e68e 100644 --- a/Documentation/mm/zsmalloc.rst +++ b/Documentation/mm/zsmalloc.rst @@ -263,3 +263,8 @@ is heavy internal fragmentation and zspool compaction is unable to relocate objects and release zspages. In these cases, it is recommended to decrease the limit on the size of the zspage chains (as specified by the CONFIG_ZSMALLOC_CHAIN_SIZE option). + +Functions +========= + +.. kernel-doc:: mm/zsmalloc.c diff --git a/Documentation/powerpc/index.rst b/Documentation/powerpc/index.rst index d33b554ca7ba..a50834798454 100644 --- a/Documentation/powerpc/index.rst +++ b/Documentation/powerpc/index.rst @@ -36,6 +36,7 @@ powerpc ultravisor vas-api vcpudispatch_stats + vmemmap_dedup features diff --git a/Documentation/powerpc/vmemmap_dedup.rst b/Documentation/powerpc/vmemmap_dedup.rst new file mode 100644 index 000000000000..dc4db59fdf87 --- /dev/null +++ b/Documentation/powerpc/vmemmap_dedup.rst @@ -0,0 +1,101 @@ +.. SPDX-License-Identifier: GPL-2.0 + +========== +Device DAX +========== + +The device-dax interface uses the tail deduplication technique explained in +Documentation/mm/vmemmap_dedup.rst + +On powerpc, vmemmap deduplication is only used with radix MMU translation. Also +with a 64K page size, only the devdax namespace with 1G alignment uses vmemmap +deduplication. + +With 2M PMD level mapping, we require 32 struct pages and a single 64K vmemmap +page can contain 1024 struct pages (64K/sizeof(struct page)). Hence there is no +vmemmap deduplication possible. + +With 1G PUD level mapping, we require 16384 struct pages and a single 64K +vmemmap page can contain 1024 struct pages (64K/sizeof(struct page)). Hence we +require 16 64K pages in vmemmap to map the struct page for 1G PUD level mapping. + +Here's how things look like on device-dax after the sections are populated:: + +-----------+ ---virt_to_page---> +-----------+ mapping to +-----------+ + | | | 0 | -------------> | 0 | + | | +-----------+ +-----------+ + | | | 1 | -------------> | 1 | + | | +-----------+ +-----------+ + | | | 2 | ----------------^ ^ ^ ^ ^ ^ + | | +-----------+ | | | | | + | | | 3 | ------------------+ | | | | + | | +-----------+ | | | | + | | | 4 | --------------------+ | | | + | PUD | +-----------+ | | | + | level | | . | ----------------------+ | | + | mapping | +-----------+ | | + | | | . | ------------------------+ | + | | +-----------+ | + | | | 15 | --------------------------+ + | | +-----------+ + | | + | | + | | + +-----------+ + + +With 4K page size, 2M PMD level mapping requires 512 struct pages and a single +4K vmemmap page contains 64 struct pages(4K/sizeof(struct page)). Hence we +require 8 4K pages in vmemmap to map the struct page for 2M pmd level mapping. + +Here's how things look like on device-dax after the sections are populated:: + + +-----------+ ---virt_to_page---> +-----------+ mapping to +-----------+ + | | | 0 | -------------> | 0 | + | | +-----------+ +-----------+ + | | | 1 | -------------> | 1 | + | | +-----------+ +-----------+ + | | | 2 | ----------------^ ^ ^ ^ ^ ^ + | | +-----------+ | | | | | + | | | 3 | ------------------+ | | | | + | | +-----------+ | | | | + | | | 4 | --------------------+ | | | + | PMD | +-----------+ | | | + | level | | 5 | ----------------------+ | | + | mapping | +-----------+ | | + | | | 6 | ------------------------+ | + | | +-----------+ | + | | | 7 | --------------------------+ + | | +-----------+ + | | + | | + | | + +-----------+ + +With 1G PUD level mapping, we require 262144 struct pages and a single 4K +vmemmap page can contain 64 struct pages (4K/sizeof(struct page)). Hence we +require 4096 4K pages in vmemmap to map the struct pages for 1G PUD level +mapping. + +Here's how things look like on device-dax after the sections are populated:: + + +-----------+ ---virt_to_page---> +-----------+ mapping to +-----------+ + | | | 0 | -------------> | 0 | + | | +-----------+ +-----------+ + | | | 1 | -------------> | 1 | + | | +-----------+ +-----------+ + | | | 2 | ----------------^ ^ ^ ^ ^ ^ + | | +-----------+ | | | | | + | | | 3 | ------------------+ | | | | + | | +-----------+ | | | | + | | | 4 | --------------------+ | | | + | PUD | +-----------+ | | | + | level | | . | ----------------------+ | | + | mapping | +-----------+ | | + | | | . | ------------------------+ | + | | +-----------+ | + | | | 4095 | --------------------------+ + | | +-----------+ + | | + | | + | | + +-----------+ diff --git a/Documentation/translations/zh_CN/mm/frontswap.rst b/Documentation/translations/zh_CN/mm/frontswap.rst deleted file mode 100644 index 434975390b48..000000000000 --- a/Documentation/translations/zh_CN/mm/frontswap.rst +++ /dev/null @@ -1,196 +0,0 @@ -:Original: Documentation/mm/frontswap.rst - -:翻译: - - 司延腾 Yanteng Si <siyanteng@loongson.cn> - -:校译: - -========= -Frontswap -========= - -Frontswap为交换页提供了一个 “transcendent memory” 的接口。在一些环境中，由 -于交换页被保存在RAM（或类似RAM的设备）中，而不是交换磁盘，因此可以获得巨大的性能 -节省（提高）。 - -.. _Transcendent memory in a nutshell: https://lwn.net/Articles/454795/ - -Frontswap之所以这么命名，是因为它可以被认为是与swap设备的“back”存储相反。存 -储器被认为是一个同步并发安全的面向页面的“伪RAM设备”，符合transcendent memory -（如Xen的“tmem”，或内核内压缩内存，又称“zcache”，或未来的类似RAM的设备）的要 -求；这个伪RAM设备不能被内核直接访问或寻址，其大小未知且可能随时间变化。驱动程序通过 -调用frontswap_register_ops将自己与frontswap链接起来，以适当地设置frontswap_ops -的功能，它提供的功能必须符合某些策略，如下所示: - -一个 “init” 将设备准备好接收与指定的交换设备编号（又称“类型”）相关的frontswap -交换页。一个 “store” 将把该页复制到transcendent memory，并与该页的类型和偏移 -量相关联。一个 “load” 将把该页，如果找到的话，从transcendent memory复制到内核 -内存，但不会从transcendent memory中删除该页。一个 “invalidate_page” 将从 -transcendent memory中删除该页，一个 “invalidate_area” 将删除所有与交换类型 -相关的页（例如，像swapoff）并通知 “device” 拒绝进一步存储该交换类型。 - -一旦一个页面被成功存储，在该页面上的匹配加载通常会成功。因此，当内核发现自己处于需 -要交换页面的情况时，它首先尝试使用frontswap。如果存储的结果是成功的，那么数据就已 -经成功的保存到了transcendent memory中，并且避免了磁盘写入，如果后来再读回数据， -也避免了磁盘读取。如果存储返回失败，transcendent memory已经拒绝了该数据，且该页 -可以像往常一样被写入交换空间。 - -请注意，如果一个页面被存储，而该页面已经存在于transcendent memory中（一个 “重复” -的存储），要么存储成功，数据被覆盖，要么存储失败，该页面被废止。这确保了旧的数据永远 -不会从frontswap中获得。 - -如果配置正确，对frontswap的监控是通过 `/sys/kernel/debug/frontswap` 目录下的 -debugfs完成的。frontswap的有效性可以通过以下方式测量（在所有交换设备中）: - -``failed_stores`` - 有多少次存储的尝试是失败的 - -``loads`` - 尝试了多少次加载（应该全部成功） - -``succ_stores`` - 有多少次存储的尝试是成功的 - -``invalidates`` - 尝试了多少次作废 - -后台实现可以提供额外的指标。 - -经常问到的问题 -============== - -* 价值在哪里? - -当一个工作负载开始交换时，性能就会下降。Frontswap通过提供一个干净的、动态的接口来 -读取和写入交换页到 “transcendent memory”，从而大大增加了许多这样的工作负载的性 -能，否则内核是无法直接寻址的。当数据被转换为不同的形式和大小（比如压缩）或者被秘密 -移动（对于一些类似RAM的设备来说，这可能对写平衡很有用）时，这个接口是理想的。交换 -页（和被驱逐的页面缓存页）是这种比RAM慢但比磁盘快得多的“伪RAM设备”的一大用途。 - -Frontswap对内核的影响相当小，为各种系统配置中更动态、更灵活的RAM利用提供了巨大的 -灵活性： - -在单一内核的情况下，又称“zcache”，页面被压缩并存储在本地内存中，从而增加了可以安 -全保存在RAM中的匿名页面总数。Zcache本质上是用压缩/解压缩的CPU周期换取更好的内存利 -用率。Benchmarks测试显示，当内存压力较低时，几乎没有影响，而在高内存压力下的一些 -工作负载上，则有明显的性能改善（25%以上）。 - -“RAMster” 在zcache的基础上增加了对集群系统的 “peer-to-peer” transcendent memory -的支持。Frontswap页面像zcache一样被本地压缩，但随后被“remotified” 到另一个系 -统的RAM。这使得RAM可以根据需要动态地来回负载平衡，也就是说，当系统A超载时，它可以 -交换到系统B，反之亦然。RAMster也可以被配置成一个内存服务器，因此集群中的许多服务器 -可以根据需要动态地交换到配置有大量内存的单一服务器上......而不需要预先配置每个客户 -有多少内存可用 - -在虚拟情况下，虚拟化的全部意义在于统计地将物理资源在多个虚拟机的不同需求之间进行复 -用。对于RAM来说，这真的很难做到，而且在不改变内核的情况下，要做好这一点的努力基本上 -是失败的（除了一些广为人知的特殊情况下的工作负载）。具体来说，Xen Transcendent Memory -后端允许管理器拥有的RAM “fallow”，不仅可以在多个虚拟机之间进行“time-shared”， -而且页面可以被压缩和重复利用，以优化RAM的利用率。当客户操作系统被诱导交出未充分利用 -的RAM时（如 “selfballooning”），突然出现的意外内存压力可能会导致交换；frontswap -允许这些页面被交换到管理器RAM中或从管理器RAM中交换（如果整体主机系统内存条件允许）， -从而减轻计划外交换可能带来的可怕的性能影响。 - -一个KVM的实现正在进行中，并且已经被RFC'ed到lkml。而且，利用frontswap，对NVM作为 -内存扩展技术的调查也在进行中。 - -* 当然，在某些情况下可能有性能上的优势，但frontswap的空间/时间开销是多少？ - -如果 CONFIG_FRONTSWAP 被禁用，每个 frontswap 钩子都会编译成空，唯一的开销是每 -个 swapon'ed swap 设备的几个额外字节。如果 CONFIG_FRONTSWAP 被启用，但没有 -frontswap的 “backend” 寄存器，每读或写一个交换页就会有一个额外的全局变量，而不 -是零。如果 CONFIG_FRONTSWAP 被启用，并且有一个frontswap的backend寄存器，并且 -后端每次 “store” 请求都失败（即尽管声称可能，但没有提供内存），CPU 的开销仍然可以 -忽略不计 - 因为每次frontswap失败都是在交换页写到磁盘之前，系统很可能是 I/O 绑定 -的，无论如何使用一小部分的 CPU 都是不相关的。 - -至于空间，如果CONFIG_FRONTSWAP被启用，并且有一个frontswap的backend注册，那么 -每个交换设备的每个交换页都会被分配一个比特。这是在内核已经为每个交换设备的每个交换 -页分配的8位（在2.6.34之前是16位）上增加的。(Hugh Dickins观察到，frontswap可能 -会偷取现有的8个比特，但是我们以后再来担心这个小的优化问题)。对于标准的4K页面大小的 -非常大的交换盘（这很罕见），这是每32GB交换盘1MB开销。 - -当交换页存储在transcendent memory中而不是写到磁盘上时，有一个副作用，即这可能会 -产生更多的内存压力，有可能超过其他的优点。一个backend，比如zcache，必须实现策略 -来仔细（但动态地）管理内存限制，以确保这种情况不会发生。 - -* 好吧，那就用内核骇客能理解的术语来快速概述一下这个frontswap补丁的作用如何？ - -我们假设在内核初始化过程中，一个frontswap 的 “backend” 已经注册了；这个注册表 -明这个frontswap 的 “backend” 可以访问一些不被内核直接访问的“内存”。它到底提 -供了多少内存是完全动态和随机的。 - -每当一个交换设备被交换时，就会调用frontswap_init()，把交换设备的编号（又称“类 -型”）作为一个参数传给它。这就通知了frontswap，以期待 “store” 与该号码相关的交 -换页的尝试。 - -每当交换子系统准备将一个页面写入交换设备时（参见swap_writepage()），就会调用 -frontswap_store。Frontswap与frontswap backend协商，如果backend说它没有空 -间，frontswap_store返回-1，内核就会照常把页换到交换设备上。注意，来自frontswap -backend的响应对内核来说是不可预测的；它可能选择从不接受一个页面，可能接受每九个 -页面，也可能接受每一个页面。但是如果backend确实接受了一个页面，那么这个页面的数 -据已经被复制并与类型和偏移量相关联了，而且backend保证了数据的持久性。在这种情况 -下，frontswap在交换设备的“frontswap_map” 中设置了一个位，对应于交换设备上的 -页面偏移量，否则它就会将数据写入该设备。 - -当交换子系统需要交换一个页面时（swap_readpage()），它首先调用frontswap_load()， -检查frontswap_map，看这个页面是否早先被frontswap backend接受。如果是，该页 -的数据就会从frontswap后端填充，换入就完成了。如果不是，正常的交换代码将被执行， -以便从真正的交换设备上获得这一页的数据。 - -所以每次frontswap backend接受一个页面时，交换设备的读取和（可能）交换设备的写 -入都被 “frontswap backend store” 和（可能）“frontswap backend loads” -所取代，这可能会快得多。 - -* frontswap不能被配置为一个 “特殊的” 交换设备，它的优先级要高于任何真正的交换 - 设备（例如像zswap，或者可能是swap-over-nbd/NFS）？ - -首先，现有的交换子系统不允许有任何种类的交换层次结构。也许它可以被重写以适应层次 -结构，但这将需要相当大的改变。即使它被重写，现有的交换子系统也使用了块I/O层，它 -假定交换设备是固定大小的，其中的任何页面都是可线性寻址的。Frontswap几乎没有触 -及现有的交换子系统，而是围绕着块I/O子系统的限制，提供了大量的灵活性和动态性。 - -例如，frontswap backend对任何交换页的接受是完全不可预测的。这对frontswap backend -的定义至关重要，因为它赋予了backend完全动态的决定权。在zcache中，人们无法预 -先知道一个页面的可压缩性如何。可压缩性 “差” 的页面会被拒绝，而 “差” 本身也可 -以根据当前的内存限制动态地定义。 - -此外，frontswap是完全同步的，而真正的交换设备，根据定义，是异步的，并且使用 -块I/O。块I/O层不仅是不必要的，而且可能进行 “优化”，这对面向RAM的设备来说是 -不合适的，包括将一些页面的写入延迟相当长的时间。同步是必须的，以确保后端的动 -态性，并避免棘手的竞争条件，这将不必要地大大增加frontswap和/或块I/O子系统的 -复杂性。也就是说，只有最初的 “store” 和 “load” 操作是需要同步的。一个独立 -的异步线程可以自由地操作由frontswap存储的页面。例如，RAMster中的 “remotification” -线程使用标准的异步内核套接字，将压缩的frontswap页面移动到远程机器。同样， -KVM的客户方实现可以进行客户内压缩，并使用 “batched” hypercalls。 - -在虚拟化环境中，动态性允许管理程序（或主机操作系统）做“intelligent overcommit”。 -例如，它可以选择只接受页面，直到主机交换可能即将发生，然后强迫客户机做他们 -自己的交换。 - -transcendent memory规格的frontswap有一个坏处。因为任何 “store” 都可 -能失败，所以必须在一个真正的交换设备上有一个真正的插槽来交换页面。因此， -frontswap必须作为每个交换设备的 “影子” 来实现，它有可能容纳交换设备可能 -容纳的每一个页面，也有可能根本不容纳任何页面。这意味着frontswap不能包含比 -swap设备总数更多的页面。例如，如果在某些安装上没有配置交换设备，frontswap -就没有用。无交换设备的便携式设备仍然可以使用frontswap，但是这种设备的 -backend必须配置某种 “ghost” 交换设备，并确保它永远不会被使用。 - - -* 为什么会有这种关于 “重复存储” 的奇怪定义？如果一个页面以前被成功地存储过， - 难道它不能总是被成功地覆盖吗？ - -几乎总是可以的，不，有时不能。考虑一个例子，数据被压缩了，原来的4K页面被压 -缩到了1K。现在，有人试图用不可压缩的数据覆盖该页，因此会占用整个4K。但是 -backend没有更多的空间了。在这种情况下，这个存储必须被拒绝。每当frontswap -拒绝一个会覆盖的存储时，它也必须使旧的数据作废，并确保它不再被访问。因为交 -换子系统会把新的数据写到读交换设备上，这是确保一致性的正确做法。 - -* 为什么frontswap补丁会创建新的头文件swapfile.h？ - -frontswap代码依赖于一些swap子系统内部的数据结构，这些数据结构多年来一直 -在静态和全局之间来回移动。这似乎是一个合理的妥协：将它们定义为全局，但在一 -个新的包含文件中声明它们，该文件不被包含swap.h的大量源文件所包含。 - -Dan Magenheimer，最后更新于2012年4月9日 diff --git a/Documentation/translations/zh_CN/mm/hugetlbfs_reserv.rst b/Documentation/translations/zh_CN/mm/hugetlbfs_reserv.rst index b7a0544224ad..0f7e7fb5ca8c 100644 --- a/Documentation/translations/zh_CN/mm/hugetlbfs_reserv.rst +++ b/Documentation/translations/zh_CN/mm/hugetlbfs_reserv.rst @@ -219,7 +219,7 @@ vma_commit_reservation()之间，预留映射有可能被改变。如果hugetlb_ 释放巨页 ======== -巨页释放是由函数free_huge_page()执行的。这个函数是hugetlbfs复合页的析构器。因此，它只传 +巨页释放是由函数free_huge_folio()执行的。这个函数是hugetlbfs复合页的析构器。因此，它只传 递一个指向页面结构体的指针。当一个巨页被释放时，可能需要进行预留计算。如果该页与包含保 留的子池相关联，或者该页在错误路径上被释放，必须恢复全局预留计数，就会出现这种情况。 @@ -387,7 +387,7 @@ region_count()在解除私有巨页映射时被调用。在私有映射中，预 然而，有几种情况是，在一个巨页被分配后，但在它被实例化之前，就遇到了错误。在这种情况下， 页面分配已经消耗了预留，并进行了适当的子池、预留映射和全局计数调整。如果页面在这个时候被释放 -（在实例化和清除PagePrivate之前），那么free_huge_page将增加全局预留计数。然而，预留映射 +（在实例化和清除PagePrivate之前），那么free_huge_folio将增加全局预留计数。然而，预留映射 显示报留被消耗了。这种不一致的状态将导致预留的巨页的 “泄漏” 。全局预留计数将比它原本的要高， 并阻止分配一个预先分配的页面。 diff --git a/Documentation/translations/zh_CN/mm/index.rst b/Documentation/translations/zh_CN/mm/index.rst index 2f53e37b8049..b950dd118be7 100644 --- a/Documentation/translations/zh_CN/mm/index.rst +++ b/Documentation/translations/zh_CN/mm/index.rst @@ -42,7 +42,6 @@ Linux内存管理文档 damon/index free_page_reporting ksm - frontswap hmm hwpoison hugetlbfs_reserv diff --git a/Documentation/translations/zh_CN/mm/split_page_table_lock.rst b/Documentation/translations/zh_CN/mm/split_page_table_lock.rst index 4fb7aa666037..a2c288670a24 100644 --- a/Documentation/translations/zh_CN/mm/split_page_table_lock.rst +++ b/Documentation/translations/zh_CN/mm/split_page_table_lock.rst @@ -56,16 +56,16 @@ Hugetlb特定的辅助函数: 架构对分页表锁的支持 ==================== -没有必要特别启用PTE分页表锁：所有需要的东西都由pgtable_pte_page_ctor() -和pgtable_pte_page_dtor()完成，它们必须在PTE表分配/释放时被调用。 +没有必要特别启用PTE分页表锁：所有需要的东西都由pagetable_pte_ctor() +和pagetable_pte_dtor()完成，它们必须在PTE表分配/释放时被调用。 确保架构不使用slab分配器来分配页表：slab使用page->slab_cache来分配其页 面。这个区域与page->ptl共享存储。 PMD分页锁只有在你有两个以上的页表级别时才有意义。 -启用PMD分页锁需要在PMD表分配时调用pgtable_pmd_page_ctor()，在释放时调 -用pgtable_pmd_page_dtor()。 +启用PMD分页锁需要在PMD表分配时调用pagetable_pmd_ctor()，在释放时调 +用pagetable_pmd_dtor()。 分配通常发生在pmd_alloc_one()中，释放发生在pmd_free()和pmd_free_tlb() 中，但要确保覆盖所有的PMD表分配/释放路径：即X86_PAE在pgd_alloc()中预先 @@ -73,7 +73,7 @@ PMD分页锁只有在你有两个以上的页表级别时才有意义。 一切就绪后，你可以设置CONFIG_ARCH_ENABLE_SPLIT_PMD_PTLOCK。 -注意：pgtable_pte_page_ctor()和pgtable_pmd_page_ctor()可能失败--必 +注意：pagetable_pte_ctor()和pagetable_pmd_ctor()可能失败--必 须正确处理。 page->ptl @@ -90,7 +90,7 @@ page->ptl用于访问分割页表锁，其中'page'是包含该表的页面struc 的指针并动态分配它。这允许在启用DEBUG_SPINLOCK或DEBUG_LOCK_ALLOC的 情况下使用分页锁，但由于间接访问而多花了一个缓存行。 -PTE表的spinlock_t分配在pgtable_pte_page_ctor()中，PMD表的spinlock_t -分配在pgtable_pmd_page_ctor()中。 +PTE表的spinlock_t分配在pagetable_pte_ctor()中，PMD表的spinlock_t +分配在pagetable_pmd_ctor()中。 请不要直接访问page->ptl - -使用适当的辅助函数。 |