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  1. 02 12月, 2020 1 次提交
    • G
      kernel: Implement selective syscall userspace redirection · 1446e1df
      由 Gabriel Krisman Bertazi 提交于 
      Introduce a mechanism to quickly disable/enable syscall handling for a
specific process and redirect to userspace via SIGSYS.  This is useful
for processes with parts that require syscall redirection and parts that
don't, but who need to perform this boundary crossing really fast,
without paying the cost of a system call to reconfigure syscall handling
on each boundary transition.  This is particularly important for Windows
games running over Wine.

The proposed interface looks like this:

  prctl(PR_SET_SYSCALL_USER_DISPATCH, <op>, <off>, <length>, [selector])

The range [<offset>,<offset>+<length>) is a part of the process memory
map that is allowed to by-pass the redirection code and dispatch
syscalls directly, such that in fast paths a process doesn't need to
disable the trap nor the kernel has to check the selector.  This is
essential to return from SIGSYS to a blocked area without triggering
another SIGSYS from rt_sigreturn.

selector is an optional pointer to a char-sized userspace memory region
that has a key switch for the mechanism. This key switch is set to
either PR_SYS_DISPATCH_ON, PR_SYS_DISPATCH_OFF to enable and disable the
redirection without calling the kernel.

The feature is meant to be set per-thread and it is disabled on
fork/clone/execv.

Internally, this doesn't add overhead to the syscall hot path, and it
requires very little per-architecture support.  I avoided using seccomp,
even though it duplicates some functionality, due to previous feedback
that maybe it shouldn't mix with seccomp since it is not a security
mechanism.  And obviously, this should never be considered a security
mechanism, since any part of the program can by-pass it by using the
syscall dispatcher.

For the sysinfo benchmark, which measures the overhead added to
executing a native syscall that doesn't require interception, the
overhead using only the direct dispatcher region to issue syscalls is
pretty much irrelevant.  The overhead of using the selector goes around
40ns for a native (unredirected) syscall in my system, and it is (as
expected) dominated by the supervisor-mode user-address access.  In
fact, with SMAP off, the overhead is consistently less than 5ns on my
test box.
Signed-off-by: NGabriel Krisman Bertazi <krisman@collabora.com>
Signed-off-by: NThomas Gleixner <tglx@linutronix.de>
Reviewed-by: NAndy Lutomirski <luto@kernel.org>
Acked-by: NPeter Zijlstra (Intel) <peterz@infradead.org>
Acked-by: NKees Cook <keescook@chromium.org>
Link: https://lore.kernel.org/r/20201127193238.821364-4-krisman@collabora.com
      1446e1df
  3. 17 10月, 2020 1 次提交
  5. 14 10月, 2020 1 次提交
  7. 07 10月, 2020 1 次提交
    • T
      x86/mce: Recover from poison found while copying from user space · c0ab7ffc
      由 Tony Luck 提交于 
      Existing kernel code can only recover from a machine check on code that
is tagged in the exception table with a fault handling recovery path.

Add two new fields in the task structure to pass information from
machine check handler to the "task_work" that is queued to run before
the task returns to user mode:

+ mce_vaddr: will be initialized to the user virtual address of the fault
  in the case where the fault occurred in the kernel copying data from
  a user address.  This is so that kill_me_maybe() can provide that
  information to the user SIGBUS handler.

+ mce_kflags: copy of the struct mce.kflags needed by kill_me_maybe()
  to determine if mce_vaddr is applicable to this error.

Add code to recover from a machine check while copying data from user
space to the kernel. Action for this case is the same as if the user
touched the poison directly; unmap the page and send a SIGBUS to the task.

Use a new helper function to share common code between the "fault
in user mode" case and the "fault while copying from user" case.

New code paths will be activated by the next patch which sets
MCE_IN_KERNEL_COPYIN.
Suggested-by: NBorislav Petkov <bp@alien8.de>
Signed-off-by: NTony Luck <tony.luck@intel.com>
Signed-off-by: NBorislav Petkov <bp@suse.de>
Link: https://lkml.kernel.org/r/20201006210910.21062-6-tony.luck@intel.com
      c0ab7ffc
  9. 03 10月, 2020 1 次提交
  11. 01 10月, 2020 1 次提交
    • J
      io_uring: don't rely on weak ->files references · 0f212204
      由 Jens Axboe 提交于 
      Grab actual references to the files_struct. To avoid circular references
issues due to this, we add a per-task note that keeps track of what
io_uring contexts a task has used. When the tasks execs or exits its
assigned files, we cancel requests based on this tracking.

With that, we can grab proper references to the files table, and no
longer need to rely on stashing away ring_fd and ring_file to check
if the ring_fd may have been closed.

Cc: stable@vger.kernel.org # v5.5+
Reviewed-by: NPavel Begunkov <asml.silence@gmail.com>
Signed-off-by: NJens Axboe <axboe@kernel.dk>
      0f212204
  13. 26 8月, 2020 2 次提交
  15. 06 8月, 2020 2 次提交
    • T
      posix-cpu-timers: Provide mechanisms to defer timer handling to task_work · 1fb497dd
      由 Thomas Gleixner 提交于 
      Running posix CPU timers in hard interrupt context has a few downsides:

 - For PREEMPT_RT it cannot work as the expiry code needs to take
   sighand lock, which is a 'sleeping spinlock' in RT. The original RT
   approach of offloading the posix CPU timer handling into a high
   priority thread was clumsy and provided no real benefit in general.

 - For fine grained accounting it's just wrong to run this in context of
   the timer interrupt because that way a process specific CPU time is
   accounted to the timer interrupt.

 - Long running timer interrupts caused by a large amount of expiring
   timers which can be created and armed by unpriviledged user space.

There is no hard requirement to expire them in interrupt context.

If the signal is targeted at the task itself then it won't be delivered
before the task returns to user space anyway. If the signal is targeted at
a supervisor process then it might be slightly delayed, but posix CPU
timers are inaccurate anyway due to the fact that they are tied to the
tick.

Provide infrastructure to schedule task work which allows splitting the
posix CPU timer code into a quick check in interrupt context and a thread
context expiry and signal delivery function. This has to be enabled by
architectures as it requires that the architecture specific KVM
implementation handles pending task work before exiting to guest mode.
Signed-off-by: NThomas Gleixner <tglx@linutronix.de>
Signed-off-by: NIngo Molnar <mingo@kernel.org>
Reviewed-by: NOleg Nesterov <oleg@redhat.com>
Acked-by: NPeter Zijlstra (Intel) <peterz@infradead.org>
Link: https://lore.kernel.org/r/20200730102337.783470146@linutronix.de
      1fb497dd
    • P
      locking/seqlock, headers: Untangle the spaghetti monster · 0cd39f46
      由 Peter Zijlstra 提交于 
      By using lockdep_assert_*() from seqlock.h, the spaghetti monster
attacked.

Attack back by reducing seqlock.h dependencies from two key high level headers:

 - <linux/seqlock.h>:               -Remove <linux/ww_mutex.h>
 - <linux/time.h>:                  -Remove <linux/seqlock.h>
 - <linux/sched.h>:                 +Add    <linux/seqlock.h>

The price was to add it to sched.h ...

Core header fallout, we add direct header dependencies instead of gaining them
parasitically from higher level headers:

 - <linux/dynamic_queue_limits.h>:  +Add <asm/bug.h>
 - <linux/hrtimer.h>:               +Add <linux/seqlock.h>
 - <linux/ktime.h>:                 +Add <asm/bug.h>
 - <linux/lockdep.h>:               +Add <linux/smp.h>
 - <linux/sched.h>:                 +Add <linux/seqlock.h>
 - <linux/videodev2.h>:             +Add <linux/kernel.h>

Arch headers fallout:

 - PARISC: <asm/timex.h>:           +Add <asm/special_insns.h>
 - SH:     <asm/io.h>:              +Add <asm/page.h>
 - SPARC:  <asm/timer_64.h>:        +Add <uapi/asm/asi.h>
 - SPARC:  <asm/vvar.h>:            +Add <asm/processor.h>, <asm/barrier.h>
                                    -Remove <linux/seqlock.h>
 - X86:    <asm/fixmap.h>:          +Add <asm/pgtable_types.h>
                                    -Remove <asm/acpi.h>

There's also a bunch of parasitic header dependency fallout in .c files, not listed
separately.

[ mingo: Extended the changelog, split up & fixed the original patch. ]
Co-developed-by: NIngo Molnar <mingo@kernel.org>
Signed-off-by: NPeter Zijlstra (Intel) <peterz@infradead.org>
Signed-off-by: NIngo Molnar <mingo@kernel.org>
Link: https://lore.kernel.org/r/20200804133438.GK2674@hirez.programming.kicks-ass.net
      0cd39f46
  17. 31 7月, 2020 2 次提交
  19. 29 7月, 2020 2 次提交
    • A
      sched: tasks: Use sequence counter with associated spinlock · b7505861
      由 Ahmed S. Darwish 提交于 
      A sequence counter write side critical section must be protected by some
form of locking to serialize writers. A plain seqcount_t does not
contain the information of which lock must be held when entering a write
side critical section.

Use the new seqcount_spinlock_t data type, which allows to associate a
spinlock with the sequence counter. This enables lockdep to verify that
the spinlock used for writer serialization is held when the write side
critical section is entered.

If lockdep is disabled this lock association is compiled out and has
neither storage size nor runtime overhead.
Signed-off-by: NAhmed S. Darwish <a.darwish@linutronix.de>
Signed-off-by: NPeter Zijlstra (Intel) <peterz@infradead.org>
Link: https://lkml.kernel.org/r/20200720155530.1173732-14-a.darwish@linutronix.de
      b7505861
    • Q
      sched/uclamp: Add a new sysctl to control RT default boost value · 13685c4a
      由 Qais Yousef 提交于 
      RT tasks by default run at the highest capacity/performance level. When
uclamp is selected this default behavior is retained by enforcing the
requested uclamp.min (p->uclamp_req[UCLAMP_MIN]) of the RT tasks to be
uclamp_none(UCLAMP_MAX), which is SCHED_CAPACITY_SCALE; the maximum
value.

This is also referred to as 'the default boost value of RT tasks'.

See commit 1a00d999 ("sched/uclamp: Set default clamps for RT tasks").

On battery powered devices, it is desired to control this default
(currently hardcoded) behavior at runtime to reduce energy consumed by
RT tasks.

For example, a mobile device manufacturer where big.LITTLE architecture
is dominant, the performance of the little cores varies across SoCs, and
on high end ones the big cores could be too power hungry.

Given the diversity of SoCs, the new knob allows manufactures to tune
the best performance/power for RT tasks for the particular hardware they
run on.

They could opt to further tune the value when the user selects
a different power saving mode or when the device is actively charging.

The runtime aspect of it further helps in creating a single kernel image
that can be run on multiple devices that require different tuning.

Keep in mind that a lot of RT tasks in the system are created by the
kernel. On Android for instance I can see over 50 RT tasks, only
a handful of which created by the Android framework.

To control the default behavior globally by system admins and device
integrator, introduce the new sysctl_sched_uclamp_util_min_rt_default
to change the default boost value of the RT tasks.

I anticipate this to be mostly in the form of modifying the init script
of a particular device.

To avoid polluting the fast path with unnecessary code, the approach
taken is to synchronously do the update by traversing all the existing
tasks in the system. This could race with a concurrent fork(), which is
dealt with by introducing sched_post_fork() function which will ensure
the racy fork will get the right update applied.

Tested on Juno-r2 in combination with the RT capacity awareness [1].
By default an RT task will go to the highest capacity CPU and run at the
maximum frequency, which is particularly energy inefficient on high end
mobile devices because the biggest core[s] are 'huge' and power hungry.

With this patch the RT task can be controlled to run anywhere by
default, and doesn't cause the frequency to be maximum all the time.
Yet any task that really needs to be boosted can easily escape this
default behavior by modifying its requested uclamp.min value
(p->uclamp_req[UCLAMP_MIN]) via sched_setattr() syscall.

[1] 804d402f: ("sched/rt: Make RT capacity-aware")
Signed-off-by: NQais Yousef <qais.yousef@arm.com>
Signed-off-by: NPeter Zijlstra (Intel) <peterz@infradead.org>
Link: https://lkml.kernel.org/r/20200716110347.19553-2-qais.yousef@arm.com
      13685c4a
  21. 28 7月, 2020 1 次提交
  23. 22 7月, 2020 1 次提交
  25. 10 7月, 2020 1 次提交
  27. 08 7月, 2020 3 次提交
  29. 04 7月, 2020 1 次提交
  31. 28 6月, 2020 2 次提交
  33. 15 6月, 2020 2 次提交
    • P
      sched: Remove sched_set_*() return value · 8b700983
      由 Peter Zijlstra 提交于 
      Ingo suggested that since the new sched_set_*() functions are
implemented using the 'nocheck' variants, they really shouldn't ever
fail, so remove the return value.

Cc: axboe@kernel.dk
Cc: daniel.lezcano@linaro.org
Cc: sudeep.holla@arm.com
Cc: airlied@redhat.com
Cc: broonie@kernel.org
Cc: paulmck@kernel.org
Suggested-by: NIngo Molnar <mingo@kernel.org>
Signed-off-by: NPeter Zijlstra (Intel) <peterz@infradead.org>
Reviewed-by: NIngo Molnar <mingo@kernel.org>
      8b700983
    • P
      sched: Provide sched_set_fifo() · 7318d4cc
      由 Peter Zijlstra 提交于 
      SCHED_FIFO (or any static priority scheduler) is a broken scheduler
model; it is fundamentally incapable of resource management, the one
thing an OS is actually supposed to do.

It is impossible to compose static priority workloads. One cannot take
two well designed and functional static priority workloads and mash
them together and still expect them to work.

Therefore it doesn't make sense to expose the priority field; the
kernel is fundamentally incapable of setting a sensible value, it
needs systems knowledge that it doesn't have.

Take away sched_setschedule() / sched_setattr() from modules and
replace them with:

  - sched_set_fifo(p); create a FIFO task (at prio 50)
  - sched_set_fifo_low(p); create a task higher than NORMAL,
	which ends up being a FIFO task at prio 1.
  - sched_set_normal(p, nice); (re)set the task to normal

This stops the proliferation of randomly chosen, and irrelevant, FIFO
priorities that dont't really mean anything anyway.

The system administrator/integrator, whoever has insight into the
actual system design and requirements (userspace) can set-up
appropriate priorities if and when needed.

Cc: airlied@redhat.com
Cc: alexander.deucher@amd.com
Cc: awalls@md.metrocast.net
Cc: axboe@kernel.dk
Cc: broonie@kernel.org
Cc: daniel.lezcano@linaro.org
Cc: gregkh@linuxfoundation.org
Cc: hannes@cmpxchg.org
Cc: herbert@gondor.apana.org.au
Cc: hverkuil@xs4all.nl
Cc: john.stultz@linaro.org
Cc: nico@fluxnic.net
Cc: paulmck@kernel.org
Cc: rafael.j.wysocki@intel.com
Cc: rmk+kernel@arm.linux.org.uk
Cc: sudeep.holla@arm.com
Cc: tglx@linutronix.de
Cc: ulf.hansson@linaro.org
Cc: wim@linux-watchdog.org
Signed-off-by: NPeter Zijlstra (Intel) <peterz@infradead.org>
Reviewed-by: NIngo Molnar <mingo@kernel.org>
Tested-by: NPaul E. McKenney <paulmck@kernel.org>
      7318d4cc
  35. 11 6月, 2020 1 次提交
    • T
      x86/{mce,mm}: Unmap the entire page if the whole page is affected and poisoned · 17fae129
      由 Tony Luck 提交于 
      An interesting thing happened when a guest Linux instance took a machine
check. The VMM unmapped the bad page from guest physical space and
passed the machine check to the guest.

Linux took all the normal actions to offline the page from the process
that was using it. But then guest Linux crashed because it said there
was a second machine check inside the kernel with this stack trace:

do_memory_failure
    set_mce_nospec
         set_memory_uc
              _set_memory_uc
                   change_page_attr_set_clr
                        cpa_flush
                             clflush_cache_range_opt

This was odd, because a CLFLUSH instruction shouldn't raise a machine
check (it isn't consuming the data). Further investigation showed that
the VMM had passed in another machine check because is appeared that the
guest was accessing the bad page.

Fix is to check the scope of the poison by checking the MCi_MISC register.
If the entire page is affected, then unmap the page. If only part of the
page is affected, then mark the page as uncacheable.

This assumes that VMMs will do the logical thing and pass in the "whole
page scope" via the MCi_MISC register (since they unmapped the entire
page).

  [ bp: Adjust to x86/entry changes. ]

Fixes: 284ce401 ("x86/memory_failure: Introduce {set, clear}_mce_nospec()")
Reported-by: NJue Wang <juew@google.com>
Signed-off-by: NTony Luck <tony.luck@intel.com>
Signed-off-by: NBorislav Petkov <bp@suse.de>
Signed-off-by: NThomas Gleixner <tglx@linutronix.de>
Tested-by: NJue Wang <juew@google.com>
Cc: <stable@vger.kernel.org>
Link: https://lkml.kernel.org/r/20200520163546.GA7977@agluck-desk2.amr.corp.intel.com
      17fae129
  37. 05 6月, 2020 1 次提交
  39. 03 6月, 2020 1 次提交
    • N
      mm/writeback: replace PF_LESS_THROTTLE with PF_LOCAL_THROTTLE · a37b0715
      由 NeilBrown 提交于 
      PF_LESS_THROTTLE exists for loop-back nfsd (and a similar need in the
loop block driver and callers of prctl(PR_SET_IO_FLUSHER)), where a
daemon needs to write to one bdi (the final bdi) in order to free up
writes queued to another bdi (the client bdi).

The daemon sets PF_LESS_THROTTLE and gets a larger allowance of dirty
pages, so that it can still dirty pages after other processses have been
throttled.  The purpose of this is to avoid deadlock that happen when
the PF_LESS_THROTTLE process must write for any dirty pages to be freed,
but it is being thottled and cannot write.

This approach was designed when all threads were blocked equally,
independently on which device they were writing to, or how fast it was.
Since that time the writeback algorithm has changed substantially with
different threads getting different allowances based on non-trivial
heuristics.  This means the simple "add 25%" heuristic is no longer
reliable.

The important issue is not that the daemon needs a *larger* dirty page
allowance, but that it needs a *private* dirty page allowance, so that
dirty pages for the "client" bdi that it is helping to clear (the bdi
for an NFS filesystem or loop block device etc) do not affect the
throttling of the daemon writing to the "final" bdi.

This patch changes the heuristic so that the task is not throttled when
the bdi it is writing to has a dirty page count below below (or equal
to) the free-run threshold for that bdi.  This ensures it will always be
able to have some pages in flight, and so will not deadlock.

In a steady-state, it is expected that PF_LOCAL_THROTTLE tasks might
still be throttled by global threshold, but that is acceptable as it is
only the deadlock state that is interesting for this flag.

This approach of "only throttle when target bdi is busy" is consistent
with the other use of PF_LESS_THROTTLE in current_may_throttle(), were
it causes attention to be focussed only on the target bdi.

So this patch
 - renames PF_LESS_THROTTLE to PF_LOCAL_THROTTLE,
 - removes the 25% bonus that that flag gives, and
 - If PF_LOCAL_THROTTLE is set, don't delay at all unless the
   global and the local free-run thresholds are exceeded.

Note that previously realtime threads were treated the same as
PF_LESS_THROTTLE threads.  This patch does *not* change the behvaiour
for real-time threads, so it is now different from the behaviour of nfsd
and loop tasks.  I don't know what is wanted for realtime.

[akpm@linux-foundation.org: coding style fixes]
Signed-off-by: NNeilBrown <neilb@suse.de>
Signed-off-by: NAndrew Morton <akpm@linux-foundation.org>
Reviewed-by: NJan Kara <jack@suse.cz>
Acked-by: Chuck Lever <chuck.lever@oracle.com>	[nfsd]
Cc: Christoph Hellwig <hch@lst.de>
Cc: Michal Hocko <mhocko@suse.com>
Cc: Trond Myklebust <trond.myklebust@hammerspace.com>
Link: http://lkml.kernel.org/r/87ftbf7gs3.fsf@notabene.neil.brown.nameSigned-off-by: NLinus Torvalds <torvalds@linux-foundation.org>
      a37b0715
  41. 28 5月, 2020 1 次提交
    • P
      sched: Replace rq::wake_list · a1488664
      由 Peter Zijlstra 提交于 
      The recent commit: 90b5363a ("sched: Clean up scheduler_ipi()")
got smp_call_function_single_async() subtly wrong. Even though it will
return -EBUSY when trying to re-use a csd, that condition is not
atomic and still requires external serialization.

The change in ttwu_queue_remote() got this wrong.

While on first reading ttwu_queue_remote() has an atomic test-and-set
that appears to serialize the use, the matching 'release' is not in
the right place to actually guarantee this serialization.

The actual race is vs the sched_ttwu_pending() call in the idle loop;
that can run the wakeup-list without consuming the CSD.

Instead of trying to chain the lists, merge them.
Signed-off-by: NPeter Zijlstra (Intel) <peterz@infradead.org>
Signed-off-by: NIngo Molnar <mingo@kernel.org>
Link: https://lore.kernel.org/r/20200526161908.129371594@infradead.org
      a1488664
  43. 19 5月, 2020 2 次提交
  45. 12 5月, 2020 1 次提交
  47. 28 4月, 2020 3 次提交
    • P
      rcu-tasks: Split ->trc_reader_need_end · 276c4104
      由 Paul E. McKenney 提交于 
      This commit splits ->trc_reader_need_end by using the rcu_special union.
This change permits readers to check to see if a memory barrier is
required without any added overhead in the common case where no such
barrier is required.  This commit also adds the read-side checking.
Later commits will add the machinery to properly set the new
->trc_reader_special.b.need_mb field.

This commit also makes rcu_read_unlock_trace_special() tolerate nested
read-side critical sections within interrupt and NMI handlers.
Signed-off-by: NPaul E. McKenney <paulmck@kernel.org>
      276c4104
    • P
      rcu-tasks: Add an RCU Tasks Trace to simplify protection of tracing hooks · d5f177d3
      由 Paul E. McKenney 提交于 
      Because RCU does not watch exception early-entry/late-exit, idle-loop,
or CPU-hotplug execution, protection of tracing and BPF operations is
needlessly complicated.  This commit therefore adds a variant of
Tasks RCU that:

o	Has explicit read-side markers to allow finite grace periods in
	the face of in-kernel loops for PREEMPT=n builds.  These markers
	are rcu_read_lock_trace() and rcu_read_unlock_trace().

o	Protects code in the idle loop, exception entry/exit, and
	CPU-hotplug code paths.  In this respect, RCU-tasks trace is
	similar to SRCU, but with lighter-weight readers.

o	Avoids expensive read-side instruction, having overhead similar
	to that of Preemptible RCU.

There are of course downsides:

o	The grace-period code can send IPIs to CPUs, even when those
	CPUs are in the idle loop or in nohz_full userspace.  This is
	mitigated by later commits.

o	It is necessary to scan the full tasklist, much as for Tasks RCU.

o	There is a single callback queue guarded by a single lock,
	again, much as for Tasks RCU.  However, those early use cases
	that request multiple grace periods in quick succession are
	expected to do so from a single task, which makes the single
	lock almost irrelevant.  If needed, multiple callback queues
	can be provided using any number of schemes.

Perhaps most important, this variant of RCU does not affect the vanilla
flavors, rcu_preempt and rcu_sched.  The fact that RCU Tasks Trace
readers can operate from idle, offline, and exception entry/exit in no
way enables rcu_preempt and rcu_sched readers to do so.

The memory ordering was outlined here:
https://lore.kernel.org/lkml/20200319034030.GX3199@paulmck-ThinkPad-P72/

This effort benefited greatly from off-list discussions of BPF
requirements with Alexei Starovoitov and Andrii Nakryiko.  At least
some of the on-list discussions are captured in the Link: tags below.
In addition, KCSAN was quite helpful in finding some early bugs.

Link: https://lore.kernel.org/lkml/20200219150744.428764577@infradead.org/
Link: https://lore.kernel.org/lkml/87mu8p797b.fsf@nanos.tec.linutronix.de/
Link: https://lore.kernel.org/lkml/20200225221305.605144982@linutronix.de/
Cc: Alexei Starovoitov <alexei.starovoitov@gmail.com>
Cc: Andrii Nakryiko <andriin@fb.com>
[ paulmck: Apply feedback from Steve Rostedt and Joel Fernandes. ]
[ paulmck: Decrement trc_n_readers_need_end upon IPI failure. ]
[ paulmck: Fix locking issue reported by rcutorture. ]
Signed-off-by: NPaul E. McKenney <paulmck@kernel.org>
      d5f177d3
    • L
      rcu: Remove unused ->rcu_read_unlock_special.b.deferred_qs field · f0bdf6d4
      由 Lai Jiangshan 提交于 
      The ->rcu_read_unlock_special.b.deferred_qs field is set to true in
rcu_read_unlock_special() but never set to false.  This is not
particularly useful, so this commit removes this field.

The only possible justification for this field is to ease debugging
of RCU deferred quiscent states, but the combination of the other
->rcu_read_unlock_special fields plus ->rcu_blocked_node and of course
->rcu_read_lock_nesting should cover debugging needs.  And if this last
proves incorrect, this patch can always be reverted, along with the
required setting of ->rcu_read_unlock_special.b.deferred_qs to false
in rcu_preempt_deferred_qs_irqrestore().
Signed-off-by: NLai Jiangshan <laijs@linux.alibaba.com>
Signed-off-by: NPaul E. McKenney <paulmck@kernel.org>
      f0bdf6d4
  49. 02 4月, 2020 1 次提交
    • E
      signal: Extend exec_id to 64bits · d1e7fd64
      由 Eric W. Biederman 提交于 
      Replace the 32bit exec_id with a 64bit exec_id to make it impossible
to wrap the exec_id counter.  With care an attacker can cause exec_id
wrap and send arbitrary signals to a newly exec'd parent.  This
bypasses the signal sending checks if the parent changes their
credentials during exec.

The severity of this problem can been seen that in my limited testing
of a 32bit exec_id it can take as little as 19s to exec 65536 times.
Which means that it can take as little as 14 days to wrap a 32bit
exec_id.  Adam Zabrocki has succeeded wrapping the self_exe_id in 7
days.  Even my slower timing is in the uptime of a typical server.
Which means self_exec_id is simply a speed bump today, and if exec
gets noticably faster self_exec_id won't even be a speed bump.

Extending self_exec_id to 64bits introduces a problem on 32bit
architectures where reading self_exec_id is no longer atomic and can
take two read instructions.  Which means that is is possible to hit
a window where the read value of exec_id does not match the written
value.  So with very lucky timing after this change this still
remains expoiltable.

I have updated the update of exec_id on exec to use WRITE_ONCE
and the read of exec_id in do_notify_parent to use READ_ONCE
to make it clear that there is no locking between these two
locations.

Link: https://lore.kernel.org/kernel-hardening/20200324215049.GA3710@pi3.com.pl
Fixes: 2.3.23pre2
Cc: stable@vger.kernel.org
Signed-off-by: N"Eric W. Biederman" <ebiederm@xmission.com>
      d1e7fd64
  51. 21 3月, 2020 2 次提交
    • S
      lockdep: Add hrtimer context tracing bits · 40db1739
      由 Sebastian Andrzej Siewior 提交于 
      Set current->irq_config = 1 for hrtimers which are not marked to expire in
hard interrupt context during hrtimer_init(). These timers will expire in
softirq context on PREEMPT_RT.

Setting this allows lockdep to differentiate these timers. If a timer is
marked to expire in hard interrupt context then the timer callback is not
supposed to acquire a regular spinlock instead of a raw_spinlock in the
expiry callback.
Signed-off-by: NSebastian Andrzej Siewior <bigeasy@linutronix.de>
Signed-off-by: NThomas Gleixner <tglx@linutronix.de>
Signed-off-by: NPeter Zijlstra (Intel) <peterz@infradead.org>
Link: https://lkml.kernel.org/r/20200321113242.534508206@linutronix.de
      40db1739
    • P
      lockdep: Introduce wait-type checks · de8f5e4f
      由 Peter Zijlstra 提交于 
      Extend lockdep to validate lock wait-type context.

The current wait-types are:

	LD_WAIT_FREE,		/* wait free, rcu etc.. */
	LD_WAIT_SPIN,		/* spin loops, raw_spinlock_t etc.. */
	LD_WAIT_CONFIG,		/* CONFIG_PREEMPT_LOCK, spinlock_t etc.. */
	LD_WAIT_SLEEP,		/* sleeping locks, mutex_t etc.. */

Where lockdep validates that the current lock (the one being acquired)
fits in the current wait-context (as generated by the held stack).

This ensures that there is no attempt to acquire mutexes while holding
spinlocks, to acquire spinlocks while holding raw_spinlocks and so on. In
other words, its a more fancy might_sleep().

Obviously RCU made the entire ordeal more complex than a simple single
value test because RCU can be acquired in (pretty much) any context and
while it presents a context to nested locks it is not the same as it
got acquired in.

Therefore its necessary to split the wait_type into two values, one
representing the acquire (outer) and one representing the nested context
(inner). For most 'normal' locks these two are the same.

[ To make static initialization easier we have the rule that:
  .outer == INV means .outer == .inner; because INV == 0. ]

It further means that its required to find the minimal .inner of the held
stack to compare against the outer of the new lock; because while 'normal'
RCU presents a CONFIG type to nested locks, if it is taken while already
holding a SPIN type it obviously doesn't relax the rules.

Below is an example output generated by the trivial test code:

  raw_spin_lock(&foo);
  spin_lock(&bar);
  spin_unlock(&bar);
  raw_spin_unlock(&foo);

 [ BUG: Invalid wait context ]
 -----------------------------
 swapper/0/1 is trying to lock:
 ffffc90000013f20 (&bar){....}-{3:3}, at: kernel_init+0xdb/0x187
 other info that might help us debug this:
 1 lock held by swapper/0/1:
  #0: ffffc90000013ee0 (&foo){+.+.}-{2:2}, at: kernel_init+0xd1/0x187

The way to read it is to look at the new -{n,m} part in the lock
description; -{3:3} for the attempted lock, and try and match that up to
the held locks, which in this case is the one: -{2,2}.

This tells that the acquiring lock requires a more relaxed environment than
presented by the lock stack.

Currently only the normal locks and RCU are converted, the rest of the
lockdep users defaults to .inner = INV which is ignored. More conversions
can be done when desired.

The check for spinlock_t nesting is not enabled by default. It's a separate
config option for now as there are known problems which are currently
addressed. The config option allows to identify these problems and to
verify that the solutions found are indeed solving them.

The config switch will be removed and the checks will permanently enabled
once the vast majority of issues has been addressed.

[ bigeasy: Move LD_WAIT_FREE,… out of CONFIG_LOCKDEP to avoid compile
	   failure with CONFIG_DEBUG_SPINLOCK + !CONFIG_LOCKDEP]
[ tglx: Add the config option ]
Signed-off-by: NPeter Zijlstra (Intel) <peterz@infradead.org>
Signed-off-by: NSebastian Andrzej Siewior <bigeasy@linutronix.de>
Signed-off-by: NThomas Gleixner <tglx@linutronix.de>
Signed-off-by: NPeter Zijlstra (Intel) <peterz@infradead.org>
Link: https://lkml.kernel.org/r/20200321113242.427089655@linutronix.de
      de8f5e4f
  53. 20 3月, 2020 1 次提交
  55. 24 2月, 2020 1 次提交
    • V
      sched/pelt: Add a new runnable average signal · 9f683953
      由 Vincent Guittot 提交于 
      Now that runnable_load_avg has been removed, we can replace it by a new
signal that will highlight the runnable pressure on a cfs_rq. This signal
track the waiting time of tasks on rq and can help to better define the
state of rqs.

At now, only util_avg is used to define the state of a rq:
  A rq with more that around 80% of utilization and more than 1 tasks is
  considered as overloaded.

But the util_avg signal of a rq can become temporaly low after that a task
migrated onto another rq which can bias the classification of the rq.

When tasks compete for the same rq, their runnable average signal will be
higher than util_avg as it will include the waiting time and we can use
this signal to better classify cfs_rqs.

The new runnable_avg will track the runnable time of a task which simply
adds the waiting time to the running time. The runnable _avg of cfs_rq
will be the /Sum of se's runnable_avg and the runnable_avg of group entity
will follow the one of the rq similarly to util_avg.
Signed-off-by: NVincent Guittot <vincent.guittot@linaro.org>
Signed-off-by: NMel Gorman <mgorman@techsingularity.net>
Signed-off-by: NIngo Molnar <mingo@kernel.org>
Reviewed-by: "Dietmar Eggemann <dietmar.eggemann@arm.com>"
Acked-by: NPeter Zijlstra <a.p.zijlstra@chello.nl>
Cc: Juri Lelli <juri.lelli@redhat.com>
Cc: Valentin Schneider <valentin.schneider@arm.com>
Cc: Phil Auld <pauld@redhat.com>
Cc: Hillf Danton <hdanton@sina.com>
Link: https://lore.kernel.org/r/20200224095223.13361-9-mgorman@techsingularity.net
      9f683953