- 02 Apr, 2020 40 commits
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Peter Xu authored
Let SH to use the new fault_signal_pending() helper. Here we'll need to move the up_read() out because that's actually needed as long as !RETRY cases. At the meantime we can drop all the rest of up_read()s now (which seems to be cleaner). Signed-off-by:
Peter Xu <peterx@redhat.com> Signed-off-by:
Andrew Morton <akpm@linux-foundation.org> Tested-by:
Brian Geffon <bgeffon@google.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Bobby Powers <bobbypowers@gmail.com> Cc: David Hildenbrand <david@redhat.com> Cc: Denis Plotnikov <dplotnikov@virtuozzo.com> Cc: "Dr . David Alan Gilbert" <dgilbert@redhat.com> Cc: Hugh Dickins <hughd@google.com> Cc: Jerome Glisse <jglisse@redhat.com> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: "Kirill A . Shutemov" <kirill@shutemov.name> Cc: Martin Cracauer <cracauer@cons.org> Cc: Marty McFadden <mcfadden8@llnl.gov> Cc: Matthew Wilcox <willy@infradead.org> Cc: Maya Gokhale <gokhale2@llnl.gov> Cc: Mel Gorman <mgorman@suse.de> Cc: Mike Kravetz <mike.kravetz@oracle.com> Cc: Mike Rapoport <rppt@linux.vnet.ibm.com> Cc: Pavel Emelyanov <xemul@openvz.org> Link: http://lkml.kernel.org/r/20200220160226.9550-1-peterx@redhat.comSigned-off-by:
Linus Torvalds <torvalds@linux-foundation.org>
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Peter Xu authored
Let powerpc code to use the new helper, by moving the signal handling earlier before the retry logic. Signed-off-by:
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Peter Xu authored
Let the arm64 fault handling to use the new fault_signal_pending() helper, by moving the signal handling out of the retry logic. Signed-off-by:
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Peter Xu authored
Let ARC to use the new helper fault_signal_pending() by moving the signal check out of the retry logic as standalone. This should also helps to simplify the code a bit. Signed-off-by:
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Peter Xu authored
Let's move the fatal signal check even earlier so that we can directly use the new fault_signal_pending() in x86 mm code. Signed-off-by:
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Peter Xu authored
For most architectures, we've got a quick path to detect fatal signal after a handle_mm_fault(). Introduce a helper for that quick path. It cleans the current codes a bit so we don't need to duplicate the same check across archs. More importantly, this will be an unified place that we handle the signal immediately right after an interrupted page fault, so it'll be much easier for us if we want to change the behavior of handling signals later on for all the archs. Note that currently only part of the archs are using this new helper, because some archs have their own way to handle signals. In the follow up patches, we'll try to apply this helper to all the rest of archs. Another note is that the "regs" parameter in the new helper is not used yet. It'll be used very soon. Now we kept it in this patch only to avoid touching all the archs again in the follow up patches. [peterx@redhat.com: fix sparse warnings] Link: http://lkml.kernel.org/r/20200311145921.GD479302@xz-x1Signed-off-by:
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Peter Xu authored
When follow_hugetlb_page() returns with *locked==0, it means we've got a VM_FAULT_RETRY within the fauling process and we've released the mmap_sem. When that happens, we should stop and bail out. Signed-off-by:
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Peter Xu authored
Patch series "mm: Page fault enhancements", v6. This series contains cleanups and enhancements to current page fault logic. The whole idea comes from the discussion between Andrea and Linus on the bug reported by syzbot here: https://lkml.org/lkml/2017/11/2/833 Basically it does two things: (a) Allows the page fault logic to be more interactive on not only SIGKILL, but also the rest of userspace signals, and, (b) Allows the page fault retry (VM_FAULT_RETRY) to happen for more than once. For (a): with the changes we should be able to react faster when page faults are working in parallel with userspace signals like SIGSTOP and SIGCONT (and more), and with that we can remove the buggy part in userfaultfd and benefit the whole page fault mechanism on faster signal processing to reach the userspace. For (b), we should be able to allow the page fault handler to loop for even more than twice. Some context: for now since we have FAULT_FLAG_ALLOW_RETRY we can allow to retry the page fault once with the same interrupt context, however never more than twice. This can be not only a potential cleanup to remove this assumption since AFAIU the code itself doesn't really have this twice-only limitation (though that should be a protective approach in the past), at the same time it'll greatly simplify future works like userfaultfd write-protect where it's possible to retry for more than twice (please have a look at [1] below for a possible user that might require the page fault to be handled for a third time; if we can remove the retry limitation we can simply drop that patch and those complexity). This patch (of 16): There's plenty of places around __get_user_pages() that has a parameter "nonblocking" which does not really mean that "it won't block" (because it can really block) but instead it shows whether the mmap_sem is released by up_read() during the page fault handling mostly when VM_FAULT_RETRY is returned. We have the correct naming in e.g. get_user_pages_locked() or get_user_pages_remote() as "locked", however there're still many places that are using the "nonblocking" as name. Renaming the places to "locked" where proper to better suite the functionality of the variable. While at it, fixing up some of the comments accordingly. Signed-off-by:
Mike Rapoport <rppt@linux.vnet.ibm.com> Reviewed-by:
Jerome Glisse <jglisse@redhat.com> Reviewed-by:
David Hildenbrand <david@redhat.com> Cc: Andrea Arcangeli <aarcange@redhat.com> Cc: Martin Cracauer <cracauer@cons.org> Cc: "Kirill A . Shutemov" <kirill@shutemov.name> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: "Dr . David Alan Gilbert" <dgilbert@redhat.com> Cc: Bobby Powers <bobbypowers@gmail.com> Cc: Maya Gokhale <gokhale2@llnl.gov> Cc: Mike Kravetz <mike.kravetz@oracle.com> Cc: Matthew Wilcox <willy@infradead.org> Cc: Marty McFadden <mcfadden8@llnl.gov> Cc: Mel Gorman <mgorman@suse.de> Cc: Hugh Dickins <hughd@google.com> Cc: Denis Plotnikov <dplotnikov@virtuozzo.com> Cc: Pavel Emelyanov <xemul@openvz.org> Link: http://lkml.kernel.org/r/20200220155353.8676-2-peterx@redhat.comSigned-off-by:
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Matthew Wilcox (Oracle) authored
The documentation currently does not include the deathless prose written to describe functions in pagemap.h because it's not included in any rst file. Fix up the mismatches between parameter names and the documentation and add the file to mm-api. Signed-off-by:
Matthew Wilcox (Oracle) <willy@infradead.org> Signed-off-by:
Zi Yan <ziy@nvidia.com> Reviewed-by:
John Hubbard <jhubbard@nvidia.com> Cc: Jonathan Corbet <corbet@lwn.net> Link: http://lkml.kernel.org/r/20200221220045.24989-1-willy@infradead.orgSigned-off-by:
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Anshuman Khandual authored
Currently the declaration and definition for is_vma_temporary_stack() are scattered. Lets make is_vma_temporary_stack() helper available for general use and also drop the declaration from (include/linux/huge_mm.h) which is no longer required. While at this, rename this as vma_is_temporary_stack() in line with existing helpers. This should not cause any functional change. Signed-off-by:
Anshuman Khandual <anshuman.khandual@arm.com> Signed-off-by:
Vlastimil Babka <vbabka@suse.cz> Cc: Ingo Molnar <mingo@redhat.com> Cc: Michael Ellerman <mpe@ellerman.id.au> Cc: Paul Mackerras <paulus@samba.org> Cc: Thomas Gleixner <tglx@linutronix.de> Link: http://lkml.kernel.org/r/1582782965-3274-4-git-send-email-anshuman.khandual@arm.comSigned-off-by:
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Anshuman Khandual authored
Idea of a foreign VMA with respect to the present context is very generic. But currently there are two identical definitions for this in powerpc and x86 platforms. Lets consolidate those redundant definitions while making vma_is_foreign() available for general use later. This should not cause any functional change. Signed-off-by:
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Anshuman Khandual authored
Patch series "mm/vma: some more minor changes", v2. The motivation here is to consolidate VMA flags and helpers in generic memory header and reduce code duplication when ever applicable. If there are other possible similar instances which might be missing here, please do let me me know. I will be happy to incorporate them. This patch (of 3): Move VM_NO_KHUGEPAGED into generic header (include/linux/mm.h). This just makes sure that no VMA flag is scattered in individual function files any longer. While at this, fix an old comment which is no longer valid. This should not cause any functional change. Signed-off-by:
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Thomas Hellstrom authored
Following the update of pagewalk code commit a07984d48146 ("mm: pagewalk: add p4d_entry() and pgd_entry()") we can modify the mapping_dirty_helpers' huge page-table entry callbacks to avoid splitting when a huge pud or -pmd is encountered. Signed-off-by:Thomas Hellstrom <thellstrom@vmware.com> Signed-off-by:
Steven Price <steven.price@arm.com> Cc: Andrew Morton <akpm@linux-foundation.org> Link: http://lkml.kernel.org/r/20200203154305.15045-1-thomas_os@shipmail.orgSigned-off-by:
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Roman Gushchin authored
If a task is getting moved out of the OOMing cgroup, it might result in unexpected OOM killings if memory.oom.group is used anywhere in the cgroup tree. Imagine the following example: A (oom.group = 1) / \ (OOM) B C Let's say B's memory.max is exceeded and it's OOMing. The OOM killer selects a task in B as a victim, but someone asynchronously moves the task into C. mem_cgroup_get_oom_group() will iterate over all ancestors of C up to the root cgroup. In theory it had to stop at the oom_domain level - the memory cgroup which is OOMing. But because B is not an ancestor of C, it's not happening. Instead it chooses A (because it's oom.group is set), and kills all tasks in A. This behavior is wrong because the OOM happened in B, so there is no reason to kill anything outside. Fix this by checking it the memory cgroup to which the task belongs is a descendant of the oom_domain. If not, memory.oom.group should be ignored, and the OOM killer should kill only the victim task. Reported-by:Dan Schatzberg <dschatzberg@fb.com> Signed-off-by:
Roman Gushchin <guro@fb.com> Signed-off-by:
Michal Hocko <mhocko@suse.com> Acked-by:
Johannes Weiner <hannes@cmpxchg.org> Link: http://lkml.kernel.org/r/20200316223510.3176148-1-guro@fb.comSigned-off-by:
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Chris Down authored
The read side of this is all protected, but we can still tear if multiple iterations of mem_cgroup_protected are going at the same time. There's some intentional racing in mem_cgroup_protected which is ok, but load/store tearing should be avoided. Signed-off-by:
Chris Down <chris@chrisdown.name> Signed-off-by:
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Chris Down authored
The write side of this is xchg()/smp_mb(), so that's all good. Just a few sites missing a READ_ONCE. Signed-off-by:
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Chris Down authored
This can be set concurrently with reads, which may cause the wrong value to be propagated. Signed-off-by:
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Chris Down authored
This can be set concurrently with reads, which may cause the wrong value to be propagated. Signed-off-by:
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Chris Down authored
This one is a bit more nuanced because we have memcg_max_mutex, which is mostly just used for enforcing invariants, but we still need to READ_ONCE since (despite its name) it doesn't really protect memory.max access. On write (page_counter_set_max() and memory_max_write()) we use xchg(), which uses smp_mb(), so that's already fine. Signed-off-by:
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Chris Down authored
A mem_cgroup's high attribute can be concurrently set at the same time as we are trying to read it -- for example, if we are in memory_high_write at the same time as we are trying to do high reclaim. Signed-off-by:
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Vincenzo Frascino authored
mem_cgroup_id_get_many() is currently used only when MMU or MEMCG_SWAP configuration options are enabled. Having them disabled triggers the following warning at compile time: linux/mm/memcontrol.c:4797:13: warning: `mem_cgroup_id_get_many' defined but not used [-Wunused-function] static void mem_cgroup_id_get_many(struct mem_cgroup *memcg, unsigned int n) Make mem_cgroup_id_get_many() __maybe_unused to address the issue. Signed-off-by:
Vincenzo Frascino <vincenzo.frascino@arm.com> Signed-off-by:
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Shakeel Butt authored
Currently multiple locations in memcg code, css_tryget_online() is being used. However it doesn't matter whether the cgroup is online for the callers. Online used to matter when we had reparenting on offlining and we needed a way to prevent new ones from showing up. The failure case for couple of these css_tryget_online usage is to fallback to root_mem_cgroup which kind of make bypassing the memcg limits possible for some workloads. For example creating an inotify group in a subcontainer and then deleting that container after moving the process to a different container will make all the event objects allocated for that group to the root_mem_cgroup. So, using css_tryget_online() is dangerous for such cases. Two locations still use the online version. The swapin of offlined memcg's pages and the memcg kmem cache creation. The kmem cache indeed needs the online version as the kernel does the reparenting of memcg kmem caches. For the swapin case, it has been left for later as the fallback is not really that concerning. With swap accounting enabled, if the memcg of the swapped out page is not online then the memcg extracted from the given 'mm' will be charged and if 'mm' is NULL then root memcg will be charged. However I could not find a code path where the given 'mm' will be NULL for swap-in case. Signed-off-by:
Shakeel Butt <shakeelb@google.com> Signed-off-by:
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Johannes Weiner authored
Right now, the effective protection of any given cgroup is capped by its own explicit memory.low setting, regardless of what the parent says. The reasons for this are mostly historical and ease of implementation: to make delegation of memory.low safe, effective protection is the min() of all memory.low up the tree. Unfortunately, this limitation makes it impossible to protect an entire subtree from another without forcing the user to make explicit protection allocations all the way to the leaf cgroups - something that is highly undesirable in real life scenarios. Consider memory in a data center host. At the cgroup top level, we have a distinction between system management software and the actual workload the system is executing. Both branches are further subdivided into individual services, job components etc. We want to protect the workload as a whole from the system management software, but that doesn't mean we want to protect and prioritize individual workload wrt each other. Their memory demand can vary over time, and we'd want the VM to simply cache the hottest data within the workload subtree. Yet, the current memory.low limitations force us to allocate a fixed amount of protection to each workload component in order to get protection from system management software in general. This results in very inefficient resource distribution. Another concern with mandating downward allocation is that, as the complexity of the cgroup tree grows, it gets harder for the lower levels to be informed about decisions made at the host-level. Consider a container inside a namespace that in turn creates its own nested tree of cgroups to run multiple workloads. It'd be extremely difficult to configure memory.low parameters in those leaf cgroups that on one hand balance pressure among siblings as the container desires, while also reflecting the host-level protection from e.g. rpm upgrades, that lie beyond one or more delegation and namespacing points in the tree. It's highly unusual from a cgroup interface POV that nested levels have to be aware of and reflect decisions made at higher levels for them to be effective. To enable such use cases and scale configurability for complex trees, this patch implements a resource inheritance model for memory that is similar to how the CPU and the IO controller implement work-conserving resource allocations: a share of a resource allocated to a subree always applies to the entire subtree recursively, while allowing, but not mandating, children to further specify distribution rules. That means that if protection is explicitly allocated among siblings, those configured shares are being followed during page reclaim just like they are now. However, if the memory.low set at a higher level is not fully claimed by the children in that subtree, the "floating" remainder is applied to each cgroup in the tree in proportion to its size. Since reclaim pressure is applied in proportion to size as well, each child in that tree gets the same boost, and the effect is neutral among siblings - with respect to each other, they behave as if no memory control was enabled at all, and the VM simply balances the memory demands optimally within the subtree. But collectively those cgroups enjoy a boost over the cgroups in neighboring trees. E.g. a leaf cgroup with a memory.low setting of 0 no longer means that it's not getting a share of the hierarchically assigned resource, just that it doesn't claim a fixed amount of it to protect from its siblings. This allows us to recursively protect one subtree (workload) from another (system management), while letting subgroups compete freely among each other - without having to assign fixed shares to each leaf, and without nested groups having to echo higher-level settings. The floating protection composes naturally with fixed protection. Consider the following example tree: A A: low = 2G / \ A1: low = 1G A1 A2 A2: low = 0G As outside pressure is applied to this tree, A1 will enjoy a fixed protection from A2 of 1G, but the remaining, unclaimed 1G from A is split evenly among A1 and A2, coming out to 1.5G and 0.5G. There is a slight risk of regressing theoretical setups where the top-level cgroups don't know about the true budgeting and set bogusly high "bypass" values that are meaningfully allocated down the tree. Such setups would rely on unclaimed protection to be discarded, and distributing it would change the intended behavior. Be safe and hide the new behavior behind a mount option, 'memory_recursiveprot'. Signed-off-by:Tejun Heo <tj@kernel.org> Acked-by:
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Johannes Weiner authored
The effective protection of any given cgroup is a somewhat complicated construct that depends on the ancestor's configuration, siblings' configurations, as well as current memory utilization in all these groups. It's done this way to satisfy hierarchical delegation requirements while also making the configuration semantics flexible and expressive in complex real life scenarios. Unfortunately, all the rules and requirements are sparsely documented, and the code is a little too clever in merging different scenarios into a single min() expression. This makes it hard to reason about the implementation and avoid breaking semantics when making changes to it. This patch documents each semantic rule individually and splits out the handling of the overcommit case from the regular case. Michal Koutný also points out that the points of equilibrium as described in the existing example scenarios aren't actually accurate. Delete these examples for now to avoid confusion. Signed-off-by:
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Johannes Weiner authored
Patch series "mm: memcontrol: recursive memory.low protection", v3. The current memory.low (and memory.min) semantics require protection to be assigned to a cgroup in an untinterrupted chain from the top-level cgroup all the way to the leaf. In practice, we want to protect entire cgroup subtrees from each other (system management software vs. workload), but we would like the VM to balance memory optimally *within* each subtree, without having to make explicit weight allocations among individual components. The current semantics make that impossible. They also introduce unmanageable complexity into more advanced resource trees. For example: host root `- system.slice `- rpm upgrades `- logging `- workload.slice `- a container `- system.slice `- workload.slice `- job A `- component 1 `- component 2 `- job B At a host-level perspective, we would like to protect the outer workload.slice subtree as a whole from rpm upgrades, logging etc. But for that to be effective, right now we'd have to propagate it down through the container, the inner workload.slice, into the job cgroup and ultimately the component cgroups where memory is actually, physically allocated. This may cross several tree delegation points and namespace boundaries, which make such a setup near impossible. CPU and IO on the other hand are already distributed recursively. The user would simply configure allowances at the host level, and they would apply to the entire subtree without any downward propagation. To enable the above-mentioned usecases and bring memory in line with other resource controllers, this patch series extends memory.low/min such that settings apply recursively to the entire subtree. Users can still assign explicit shares in subgroups, but if they don't, any ancestral protection will be distributed such that children compete freely amongst each other - as if no memory control were enabled inside the subtree - but enjoy protection from neighboring trees. In the above example, the user would then be able to configure shares of CPU, IO and memory at the host level to comprehensively protect and isolate the workload.slice as a whole from system.slice activity. Patch #1 fixes an existing bug that can give a cgroup tree more protection than it should receive as per ancestor configuration. Patch #2 simplifies and documents the existing code to make it easier to reason about the changes in the next patch. Patch #3 finally implements recursive memory protection semantics. Because of a risk of regressing legacy setups, the new semantics are hidden behind a cgroup2 mount option, 'memory_recursiveprot'. More details in patch #3. This patch (of 3): When memory.low is overcommitted - i.e. the children claim more protection than their shared ancestor grants them - the allowance is distributed in proportion to how much each sibling uses their own declared protection: low_usage = min(memory.low, memory.current) elow = parent_elow * (low_usage / siblings_low_usage) However, siblings_low_usage is not the sum of all low_usages. It sums up the usages of *only those cgroups that are within their memory.low* That means that low_usage can be *bigger* than siblings_low_usage, and consequently the total protection afforded to the children can be bigger than what the ancestor grants the subtree. Consider three groups where two are in excess of their protection: A/memory.low = 10G A/A1/memory.low = 10G, memory.current = 20G A/A2/memory.low = 10G, memory.current = 20G A/A3/memory.low = 10G, memory.current = 8G siblings_low_usage = 8G (only A3 contributes) A1/elow = parent_elow(10G) * low_usage(10G) / siblings_low_usage(8G) = 12.5G -> 10G A2/elow = parent_elow(10G) * low_usage(10G) / siblings_low_usage(8G) = 12.5G -> 10G A3/elow = parent_elow(10G) * low_usage(8G) / siblings_low_usage(8G) = 10.0G (the 12.5G are capped to the explicit memory.low setting of 10G) With that, the sum of all awarded protection below A is 30G, when A only grants 10G for the entire subtree. What does this mean in practice? A1 and A2 would still be in excess of their 10G allowance and would be reclaimed, whereas A3 would not. As they eventually drop below their protection setting, they would be counted in siblings_low_usage again and the error would right itself. When reclaim was applied in a binary fashion (cgroup is reclaimed when it's above its protection, otherwise it's skipped) this would actually work out just fine. However, since 1bc63fb1 ("mm, memcg: make scan aggression always exclude protection"), reclaim pressure is scaled to how much a cgroup is above its protection. As a result this calculation error unduly skews pressure away from A1 and A2 toward the rest of the system. But why did we do it like this in the first place? The reasoning behind exempting groups in excess from siblings_low_usage was to go after them first during reclaim in an overcommitted subtree: A/memory.low = 2G, memory.current = 4G A/A1/memory.low = 3G, memory.current = 2G A/A2/memory.low = 1G, memory.current = 2G siblings_low_usage = 2G (only A1 contributes) A1/elow = parent_elow(2G) * low_usage(2G) / siblings_low_usage(2G) = 2G A2/elow = parent_elow(2G) * low_usage(1G) / siblings_low_usage(2G) = 1G While the children combined are overcomitting A and are technically both at fault, A2 is actively declaring unprotected memory and we would like to reclaim that first. However, while this sounds like a noble goal on the face of it, it doesn't make much difference in actual memory distribution: Because A is overcommitted, reclaim will not stop once A2 gets pushed back to within its allowance; we'll have to reclaim A1 either way. The end result is still that protection is distributed proportionally, with A1 getting 3/4 (1.5G) and A2 getting 1/4 (0.5G) of A's allowance. [ If A weren't overcommitted, it wouldn't make a difference since each cgroup would just get the protection it declares: A/memory.low = 2G, memory.current = 3G A/A1/memory.low = 1G, memory.current = 1G A/A2/memory.low = 1G, memory.current = 2G With the current calculation: siblings_low_usage = 1G (only A1 contributes) A1/elow = parent_elow(2G) * low_usage(1G) / siblings_low_usage(1G) = 2G -> 1G A2/elow = parent_elow(2G) * low_usage(1G) / siblings_low_usage(1G) = 2G -> 1G Including excess groups in siblings_low_usage: siblings_low_usage = 2G A1/elow = parent_elow(2G) * low_usage(1G) / siblings_low_usage(2G) = 1G -> 1G A2/elow = parent_elow(2G) * low_usage(1G) / siblings_low_usage(2G) = 1G -> 1G ] Simplify the calculation and fix the proportional reclaim bug by including excess cgroups in siblings_low_usage. After this patch, the effective memory.low distribution from the example above would be as follows: A/memory.low = 10G A/A1/memory.low = 10G, memory.current = 20G A/A2/memory.low = 10G, memory.current = 20G A/A3/memory.low = 10G, memory.current = 8G siblings_low_usage = 28G A1/elow = parent_elow(10G) * low_usage(10G) / siblings_low_usage(28G) = 3.5G A2/elow = parent_elow(10G) * low_usage(10G) / siblings_low_usage(28G) = 3.5G A3/elow = parent_elow(10G) * low_usage(8G) / siblings_low_usage(28G) = 2.8G Fixes: 1bc63fb1 ("mm, memcg: make scan aggression always exclude protection") Fixes: 23067153 ("mm: memory.low hierarchical behavior") Signed-off-by: -
Roman Gushchin authored
Drop the _memcg suffix from (__)memcg_kmem_(un)charge functions. It's shorter and more obvious. These are the most basic functions which are just (un)charging the given cgroup with the given amount of pages. Also fix up the corresponding comments. Signed-off-by:
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Roman Gushchin authored
There are many places in memcg_charge_slab() and memcg_uncharge_slab() which are calculating the number of pages to charge, css references to grab etc depending on the order of the slab page. Let's simplify the code by calculating it once and caching in the local variable. Signed-off-by:
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Roman Gushchin authored
These functions are charging the given number of kernel pages to the given memory cgroup. The number doesn't have to be a power of two. Let's make them to take the unsigned int nr_pages as an argument instead of the page order. It makes them look consistent with the corresponding uncharge functions and functions like: mem_cgroup_charge_skmem(memcg, nr_pages). Signed-off-by:
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Roman Gushchin authored
Rename (__)memcg_kmem_(un)charge() into (__)memcg_kmem_(un)charge_page() to better reflect what they are actually doing: 1) call __memcg_kmem_(un)charge_memcg() to actually charge or uncharge the current memcg 2) set or clear the PageKmemcg flag Signed-off-by:
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Roman Gushchin authored
Drop the unused page argument and put the memcg pointer at the first place. This make the function consistent with its peers: __memcg_kmem_uncharge_memcg(), memcg_kmem_charge_memcg(), etc. Signed-off-by:
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Roman Gushchin authored
Patch series "mm: memcg: kmem API cleanup", v2. This patchset aims to clean up the kernel memory charging API. It doesn't bring any functional changes, just removes unused arguments, renames some functions and fixes some comments. Currently it's not obvious which functions are most basic (memcg_kmem_(un)charge_memcg()) and which are based on them (memcg_kmem_(un)charge()). The patchset renames these functions and removes unused arguments: TL;DR: was: memcg_kmem_charge_memcg(page, gfp, order, memcg) memcg_kmem_uncharge_memcg(memcg, nr_pages) memcg_kmem_charge(page, gfp, order) memcg_kmem_uncharge(page, order) now: memcg_kmem_charge(memcg, gfp, nr_pages) memcg_kmem_uncharge(memcg, nr_pages) memcg_kmem_charge_page(page, gfp, order) memcg_kmem_uncharge_page(page, order) This patch (of 6): The first argument of memcg_kmem_charge_memcg() and __memcg_kmem_charge_memcg() is the page pointer and it's not used. Let's drop it. Memcg pointer is passed as the last argument. Move it to the first place for consistency with other memcg functions, e.g. __memcg_kmem_uncharge_memcg() or try_charge(). Signed-off-by:
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Roman Gushchin authored
Sometimes we need to get a memcg pointer from a charged kernel object. The right way to get it depends on whether it's a proper slab object or it's backed by raw pages (e.g. it's a vmalloc alloction). In the first case the kmem_cache->memcg_params.memcg indirection should be used; in other cases it's just page->mem_cgroup. To simplify this task and hide the implementation details let's use the mem_cgroup_from_obj() helper, which takes a pointer to any kernel object and returns a valid memcg pointer or NULL. Passing a kernel address rather than a pointer to a page will allow to use this helper for per-object (rather than per-page) tracked objects in the future. The caller is still responsible to ensure that the returned memcg isn't going away underneath: take the rcu read lock, cgroup mutex etc; depending on the context. mem_cgroup_from_kmem() defined in mm/list_lru.c is now obsolete and can be removed. Signed-off-by:
Yafang Shao <laoar.shao@gmail.com> Cc: Michal Hocko <mhocko@kernel.org> Cc: Johannes Weiner <hannes@cmpxchg.org> Cc: Vladimir Davydov <vdavydov.dev@gmail.com> Link: http://lkml.kernel.org/r/20200117203609.3146239-1-guro@fb.comSigned-off-by:
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Kirill Tkhai authored
The shrinker_map may be touched from any cpu (e.g., a bit there may be set by a task running everywhere) but kswapd is always bound to specific node. So allocate shrinker_map from the related NUMA node to respect its NUMA locality. Also, this follows generic way we use for allocation of memcg's per-node data. Signed-off-by:
Kirill Tkhai <ktkhai@virtuozzo.com> Signed-off-by:
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Yafang Shao authored
When I manually set default n to MEMCG_KMEM in init/Kconfig, bellow error occurs, mm/slab_common.c: In function 'memcg_slab_start': mm/slab_common.c:1530:30: error: 'struct mem_cgroup' has no member named 'kmem_caches' return seq_list_start(&memcg->kmem_caches, *pos); ^ mm/slab_common.c: In function 'memcg_slab_next': mm/slab_common.c:1537:32: error: 'struct mem_cgroup' has no member named 'kmem_caches' return seq_list_next(p, &memcg->kmem_caches, pos); ^ mm/slab_common.c: In function 'memcg_slab_show': mm/slab_common.c:1551:16: error: 'struct mem_cgroup' has no member named 'kmem_caches' if (p == memcg->kmem_caches.next) ^ CC arch/x86/xen/smp.o mm/slab_common.c: In function 'memcg_slab_start': mm/slab_common.c:1531:1: warning: control reaches end of non-void function [-Wreturn-type] } ^ mm/slab_common.c: In function 'memcg_slab_next': mm/slab_common.c:1538:1: warning: control reaches end of non-void function [-Wreturn-type] } ^ That's because kmem_caches is defined only when CONFIG_MEMCG_KMEM is set, while memcg_slab_start() will use it no matter CONFIG_MEMCG_KMEM is defined or not. By the way, the reason I mannuly undefined CONFIG_MEMCG_KMEM is to verify whether my some other code change is still stable when CONFIG_MEMCG_KMEM is not set. Unfortunately, the existing code has been already unstable since v4.11. Fixes: bc2791f8 ("slab: link memcg kmem_caches on their associated memory cgroup") Signed-off-by: -
Wei Yang authored
add_to_swap_cache() and delete_from_swap_cache() are counterparts, while currently they use different ways to count pages. It doesn't break anything because we only have two sizes for PageAnon, but this is confusing and not good practice. This patch corrects it by making both functions use hpage_nr_pages(). Signed-off-by:
Wei Yang <richard.weiyang@gmail.com> Signed-off-by:
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Yang Shi authored
Memory barrier is needed after setting LRU bit, but smp_mb() is too strong. Some architectures, i.e. x86, imply memory barrier with atomic operations, so replacing it with smp_mb__after_atomic() sounds better, which is nop on strong ordered machines, and full memory barriers on others. With this change the vm-scalability cases would perform better on x86, I saw total 6% improvement with this patch and previous inline fix. The test data (lru-file-readtwice throughput) against v5.6-rc4: mainline w/ inline fix w/ both (adding this) 150MB 154MB 159MB Fixes: 9c4e6b1a ("mm, mlock, vmscan: no more skipping pagevecs") Signed-off-by:
Yang Shi <yang.shi@linux.alibaba.com> Signed-off-by:
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Yang Shi authored
When backporting commit 9c4e6b1a ("mm, mlock, vmscan: no more skipping pagevecs") to our 4.9 kernel, our test bench noticed around 10% down with a couple of vm-scalability's test cases (lru-file-readonce, lru-file-readtwice and lru-file-mmap-read). I didn't see that much down on my VM (32c-64g-2nodes). It might be caused by the test configuration, which is 32c-256g with NUMA disabled and the tests were run in root memcg, so the tests actually stress only one inactive and active lru. It sounds not very usual in mordern production environment. That commit did two major changes: 1. Call page_evictable() 2. Use smp_mb to force the PG_lru set visible It looks they contribute the most overhead. The page_evictable() is a function which does function prologue and epilogue, and that was used by page reclaim path only. However, lru add is a very hot path, so it sounds better to make it inline. However, it calls page_mapping() which is not inlined either, but the disassemble shows it doesn't do push and pop operations and it sounds not very straightforward to inline it. Other than this, it sounds smp_mb() is not necessary for x86 since SetPageLRU is atomic which enforces memory barrier already, replace it with smp_mb__after_atomic() in the following patch. With the two fixes applied, the tests can get back around 5% on that test bench and get back normal on my VM. Since the test bench configuration is not that usual and I also saw around 6% up on the latest upstream, so it sounds good enough IMHO. The below is test data (lru-file-readtwice throughput) against the v5.6-rc4: mainline w/ inline fix 150MB 154MB With this patch the throughput gets 2.67% up. The data with using smp_mb__after_atomic() is showed in the following patch. Shakeel Butt did the below test: On a real machine with limiting the 'dd' on a single node and reading 100 GiB sparse file (less than a single node). Just ran a single instance to not cause the lru lock contention. The cmdline used is "dd if=file-100GiB of=/dev/null bs=4k". Ran the cmd 10 times with drop_caches in between and measured the time it took. Without patch: 56.64143 +- 0.672 sec With patches: 56.10 +- 0.21 sec [akpm@linux-foundation.org: move page_evictable() to internal.h] Fixes: 9c4e6b1a ("mm, mlock, vmscan: no more skipping pagevecs") Signed-off-by:
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Wei Yang authored
Currently we use a tmp pointer, pentry, to transfer and reset swap cache slot, which is a little redundant. Swap cache slot stores the entry value directly, assign and reset it by value would be straight forward. Also this patch merges the else and if, since this is the only case we refill and repeat swap cache. Signed-off-by:
Wei Yang <richard.weiyang@linux.alibaba.com> Signed-off-by:
Tim Chen <tim.c.chen@linux.intel.com> Link: http://lkml.kernel.org/r/20200311055352.50574-1-richard.weiyang@linux.alibaba.comSigned-off-by:
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Qian Cai authored
si->inuse_pages could be accessed concurrently as noticed by KCSAN, write to 0xffff98b00ebd04dc of 4 bytes by task 82262 on cpu 92: swap_range_free+0xbe/0x230 swap_range_free at mm/swapfile.c:719 swapcache_free_entries+0x1be/0x250 free_swap_slot+0x1c8/0x220 __swap_entry_free.constprop.19+0xa3/0xb0 free_swap_and_cache+0x53/0xa0 unmap_page_range+0x7e0/0x1ce0 unmap_single_vma+0xcd/0x170 unmap_vmas+0x18b/0x220 exit_mmap+0xee/0x220 mmput+0xe7/0x240 do_exit+0x598/0xfd0 do_group_exit+0x8b/0x180 get_signal+0x293/0x13d0 do_signal+0x37/0x5d0 prepare_exit_to_usermode+0x1b7/0x2c0 ret_from_intr+0x32/0x42 read to 0xffff98b00ebd04dc of 4 bytes by task 82499 on cpu 46: try_to_unuse+0x86b/0xc80 try_to_unuse at mm/swapfile.c:2185 __x64_sys_swapoff+0x372/0xd40 do_syscall_64+0x91/0xb05 entry_SYSCALL_64_after_hwframe+0x49/0xbe The plain reads in try_to_unuse() are outside si->lock critical section which result in data races that could be dangerous to be used in a loop. Fix them by adding READ_ONCE(). Signed-off-by:
Qian Cai <cai@lca.pw> Signed-off-by:
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Wei Yang authored
__pagevec_lru_add() is only used in mm directory now. Remove the export symbol. Signed-off-by:
Wei Yang <richardw.yang@linux.intel.com> Signed-off-by:
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