- 04 Nov, 2014 14 commits
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Wanpeng Li authored
This patch pushes task away if the dealine of the task is equal to current during wake up. The same behavior as rt class. Signed-off-by:
Wanpeng Li <wanpeng.li@linux.intel.com> Signed-off-by:
Peter Zijlstra (Intel) <peterz@infradead.org> Cc: Juri Lelli <juri.lelli@arm.com> Cc: Kirill Tkhai <ktkhai@parallels.com> Cc: Linus Torvalds <torvalds@linux-foundation.org> Link: http://lkml.kernel.org/r/1414708776-124078-4-git-send-email-wanpeng.li@linux.intel.comSigned-off-by:
Ingo Molnar <mingo@kernel.org>
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Wanpeng Li authored
This patch add deadline rq status print. Signed-off-by:
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Wanpeng Li authored
The yield semantic of deadline class is to reduce remaining runtime to zero, and then update_curr_dl() will stop it. However, comsumed bandwidth is reduced from the budget of yield task again even if it has already been set to zero which leads to artificial overrun. This patch fix it by make sure we don't steal some more time from the task that yielded in update_curr_dl(). Suggested-by:
Juri Lelli <juri.lelli@arm.com> Signed-off-by:
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Wanpeng Li authored
This patch checks if current can be pushed/pulled somewhere else in advance to make logic clear, the same behavior as dl class. - If current can't be migrated, useless to reschedule, let's hope task can move out. - If task is migratable, so let's not schedule it and see if it can be pushed or pulled somewhere else. Signed-off-by:
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Juri Lelli authored
As per commit f10e00f4 ("sched/dl: Use dl_bw_of() under rcu_read_lock_sched()"), dl_bw_of() has to be protected by rcu_read_lock_sched(). Signed-off-by:
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Yao Dongdong authored
Idle cpu is idler than non-idle cpu, so we needn't search for least_loaded_cpu after we have found an idle cpu. Signed-off-by:
Yao Dongdong <yaodongdong@huawei.com> Reviewed-by:
Srikar Dronamraju <srikar@linux.vnet.ibm.com> Signed-off-by:
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Kirill Tkhai authored
Currently used hrtimer_try_to_cancel() is racy: raw_spin_lock(&rq->lock) ... dl_task_timer raw_spin_lock(&rq->lock) ... raw_spin_lock(&rq->lock) ... switched_from_dl() ... ... hrtimer_try_to_cancel() ... ... switched_to_fair() ... ... ... ... ... ... ... ... raw_spin_unlock(&rq->lock) ... (asquired) ... ... ... ... ... ... do_exit() ... ... schedule() ... ... raw_spin_lock(&rq->lock) ... raw_spin_unlock(&rq->lock) ... ... ... raw_spin_unlock(&rq->lock) ... raw_spin_lock(&rq->lock) ... ... (asquired) put_task_struct() ... ... free_task_struct() ... ... ... ... raw_spin_unlock(&rq->lock) ... (asquired) ... ... ... ... ... (use after free) ... So, let's implement 100% guaranteed way to cancel the timer and let's be sure we are safe even in very unlikely situations. rq unlocking does not limit the area of switched_from_dl() use, because this has already been possible in pull_dl_task() below. Let's consider the safety of of this unlocking. New code in the patch is working when hrtimer_try_to_cancel() fails. This means the callback is running. In this case hrtimer_cancel() is just waiting till the callback is finished. Two 1) Since we are in switched_from_dl(), new class is not dl_sched_class and new prio is not less MAX_DL_PRIO. So, the callback returns early; it's right after !dl_task() check. After that hrtimer_cancel() returns back too. The above is: raw_spin_lock(rq->lock); ... ... dl_task_timer() ... raw_spin_lock(rq->lock); switched_from_dl() ... hrtimer_try_to_cancel() ... raw_spin_unlock(rq->lock); ... hrtimer_cancel() ... ... raw_spin_unlock(rq->lock); ... return HRTIMER_NORESTART; ... ... raw_spin_lock(rq->lock); ... 2) But the below is also possible: dl_task_timer() raw_spin_lock(rq->lock); ... raw_spin_unlock(rq->lock); raw_spin_lock(rq->lock); ... switched_from_dl() ... hrtimer_try_to_cancel() ... ... return HRTIMER_NORESTART; raw_spin_unlock(rq->lock); ... hrtimer_cancel(); ... raw_spin_lock(rq->lock); ... In this case hrtimer_cancel() returns immediately. Very unlikely case, just to mention. Nobody can manipulate the task, because check_class_changed() is always called with pi_lock locked. Nobody can force the task to participate in (concurrent) priority inheritance schemes (the same reason). All concurrent task operations require pi_lock, which is held by us. No deadlocks with dl_task_timer() are possible, because it returns right after !dl_task() check (it does nothing). If we receive a new dl_task during the time of unlocked rq, we just don't have to do pull_dl_task() in switched_from_dl() further. Signed-off-by:Kirill Tkhai <ktkhai@parallels.com> [ Added comments] Signed-off-by:
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Peter Zijlstra authored
In some cases this can trigger a true flood of output. Requested-by:
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Peter Zijlstra authored
rtnl_lock_unregistering*() take rtnl_lock() -- a mutex -- inside a wait loop. The wait loop relies on current->state to function, but so does mutex_lock(), nesting them makes for the inner to destroy the outer state. Fix this using the new wait_woken() bits. Reported-by:
Fengguang Wu <fengguang.wu@intel.com> Signed-off-by:
David S. Miller <davem@davemloft.net> Cc: Oleg Nesterov <oleg@redhat.com> Cc: Cong Wang <cwang@twopensource.com> Cc: David Gibson <david@gibson.dropbear.id.au> Cc: Eric Biederman <ebiederm@xmission.com> Cc: Eric Dumazet <edumazet@google.com> Cc: Jamal Hadi Salim <jhs@mojatatu.com> Cc: Jerry Chu <hkchu@google.com> Cc: Jiri Pirko <jiri@resnulli.us> Cc: John Fastabend <john.fastabend@gmail.com> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: Nicolas Dichtel <nicolas.dichtel@6wind.com> Cc: sfeldma@cumulusnetworks.com <sfeldma@cumulusnetworks.com> Cc: stephen hemminger <stephen@networkplumber.org> Cc: Tom Gundersen <teg@jklm.no> Cc: Tom Herbert <therbert@google.com> Cc: Veaceslav Falico <vfalico@gmail.com> Cc: Vlad Yasevich <vyasevic@redhat.com> Cc: netdev@vger.kernel.org Link: http://lkml.kernel.org/r/20141029173110.GE15602@worktop.programming.kicks-ass.netSigned-off-by:
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Peter Zijlstra authored
rfcomm_run() is a tad broken in that is has a nested wait loop. One cannot rely on p->state for the outer wait because the inner wait will overwrite it. Fix this using the new wait_woken() facility. Signed-off-by:
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Peter Zijlstra authored
The kauditd_thread wait loop is a bit iffy; it has a number of problems: - calls try_to_freeze() before schedule(); you typically want the thread to re-evaluate the sleep condition when unfreezing, also freeze_task() issues a wakeup. - it unconditionally does the {add,remove}_wait_queue(), even when the sleep condition is false. Use wait_event_freezable() that does the right thing. Reported-by:Mike Galbraith <umgwanakikbuti@gmail.com> Signed-off-by:
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Peter Zijlstra (Intel) authored
There is no user.. make it go away. Signed-off-by:
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Peter Zijlstra authored
Provide better implementations of wait_event_freezable() APIs. The problem is with freezer_do_not_count(), it hides the thread from the freezer, even though this thread might not actually freeze/sleep at all. Cc: oleg@redhat.com Cc: Rafael Wysocki <rjw@rjwysocki.net> Signed-off-by:
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Peter Zijlstra authored
There is a race between kthread_stop() and the new wait_woken() that can result in a lack of progress. CPU 0 | CPU 1 | rfcomm_run() | kthread_stop() ... | if (!test_bit(KTHREAD_SHOULD_STOP)) | | set_bit(KTHREAD_SHOULD_STOP) | wake_up_process() wait_woken() | wait_for_completion() set_current_state(INTERRUPTIBLE) | if (!WQ_FLAG_WOKEN) | schedule_timeout() | | After which both tasks will wait.. forever. Fix this by having wait_woken() check for kthread_should_stop() but only for kthreads (obviously). Signed-off-by:
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- 28 Oct, 2014 26 commits
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Peter Zijlstra authored
cond_resched() is a preemption point, not strictly a blocking primitive, so exclude it from the ->state test. In particular, preemption preserves task_struct::state. Signed-off-by:
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Peter Zijlstra authored
Validate we call might_sleep() with TASK_RUNNING, which catches places where we nest blocking primitives, eg. mutex usage in a wait loop. Since all blocking is arranged through task_struct::state, nesting this will cause the inner primitive to set TASK_RUNNING and the outer will thus not block. Another observed problem is calling a blocking function from schedule()->sched_submit_work()->blk_schedule_flush_plug() which will then destroy the task state for the actual __schedule() call that comes after it. Signed-off-by:
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Peter Zijlstra authored
WARNING: CPU: 1 PID: 1744 at kernel/sched/core.c:7104 __might_sleep+0x58/0x90() do not call blocking ops when !TASK_RUNNING; state=1 set at [<ffffffff81070e10>] prepare_to_wait+0x50 /0xa0 [<ffffffff8105bc38>] __might_sleep+0x58/0x90 [<ffffffff8148c671>] lock_sock_nested+0x31/0xb0 [<ffffffff81498aaa>] sk_stream_wait_memory+0x18a/0x2d0 Which is a false positive because sk_wait_event() will already have TASK_RUNNING at that point if it would've gone through schedule_timeout(). So annotate with sched_annotate_sleep(); which goes away on !DEBUG builds. Reported-by:
Ilya Dryomov <ilya.dryomov@inktank.com> Signed-off-by:
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Peter Zijlstra authored
This is a genuine bug in add_unformed_module(), we cannot use blocking primitives inside a wait loop. So rewrite the wait_event_interruptible() usage to use the fresh wait_woken() stuff. Reported-by:
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Peter Zijlstra authored
smp_hotplug_thread::{setup,unpark} functions can sleep too, so be consistent and do the same for all callbacks. __might_sleep+0x74/0x80 kmem_cache_alloc_trace+0x4e/0x1c0 perf_event_alloc+0x55/0x450 perf_event_create_kernel_counter+0x2f/0x100 watchdog_nmi_enable+0x8d/0x160 watchdog_enable+0x45/0x90 smpboot_thread_fn+0xec/0x2b0 kthread+0xe4/0x100 ret_from_fork+0x7c/0xb0 Signed-off-by: -
Peter Zijlstra authored
n_tty_{read,write} are wait loops with sleeps in. Wait loops rely on task_struct::state and sleeps do too, since that's the only means of actually sleeping. Therefore the nested sleeps destroy the wait loop state. Fix this by using the new woken_wake_function and wait_woken() stuff, which registers wakeups in wait and thereby allows shrinking the task_state::state changes to the actual sleep part. Signed-off-by:Greg Kroah-Hartman <gregkh@linuxfoundation.org> Cc: Jiri Slaby <jslaby@suse.cz> Cc: Linus Torvalds <torvalds@linux-foundation.org> Cc: tglx@linutronix.de Cc: ilya.dryomov@inktank.com Cc: umgwanakikbuti@gmail.com Cc: oleg@redhat.com Link: http://lkml.kernel.org/r/20140924082242.323011233@infradead.orgSigned-off-by:
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Peter Zijlstra authored
inotify_read is a wait loop with sleeps in. Wait loops rely on task_struct::state and sleeps do too, since that's the only means of actually sleeping. Therefore the nested sleeps destroy the wait loop state and the wait loop breaks the sleep functions that assume TASK_RUNNING (mutex_lock). Fix this by using the new woken_wake_function and wait_woken() stuff, which registers wakeups in wait and thereby allows shrinking the task_state::state changes to the actual sleep part. Signed-off-by:
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Peter Zijlstra authored
do_wait() is a big wait loop, but we set TASK_RUNNING too late; we end up calling potential sleeps before we reset it. Not strictly a bug since we're guaranteed to exit the loop and not call schedule(); put in annotations to quiet might_sleep(). WARNING: CPU: 0 PID: 1 at ../kernel/sched/core.c:7123 __might_sleep+0x7e/0x90() do not call blocking ops when !TASK_RUNNING; state=1 set at [<ffffffff8109a788>] do_wait+0x88/0x270 Call Trace: [<ffffffff81694991>] dump_stack+0x4e/0x7a [<ffffffff8109877c>] warn_slowpath_common+0x8c/0xc0 [<ffffffff8109886c>] warn_slowpath_fmt+0x4c/0x50 [<ffffffff810bca6e>] __might_sleep+0x7e/0x90 [<ffffffff811a1c15>] might_fault+0x55/0xb0 [<ffffffff8109a3fb>] wait_consider_task+0x90b/0xc10 [<ffffffff8109a804>] do_wait+0x104/0x270 [<ffffffff8109b837>] SyS_wait4+0x77/0x100 [<ffffffff8169d692>] system_call_fastpath+0x16/0x1b Signed-off-by:
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Peter Zijlstra authored
Add more might_sleep() checks, suppose someone put a wait_event() like thing in a wait loop.. Can't put might_sleep() in ___wait_event() because there's the locked primitives which call ___wait_event() with locks held. Signed-off-by:
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Peter Zijlstra authored
There are a few places that call blocking primitives from wait loops, provide infrastructure to support this without the typical task_struct::state collision. We record the wakeup in wait_queue_t::flags which leaves task_struct::state free to be used by others. Signed-off-by:
Oleg Nesterov <oleg@redhat.com> Cc: tglx@linutronix.de Cc: ilya.dryomov@inktank.com Cc: umgwanakikbuti@gmail.com Cc: Linus Torvalds <torvalds@linux-foundation.org> Link: http://lkml.kernel.org/r/20140924082242.051202318@infradead.orgSigned-off-by:
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Peter Zijlstra authored
We're going to make might_sleep() test for TASK_RUNNING, because blocking without TASK_RUNNING will destroy the task state by setting it to TASK_RUNNING. There are a few occasions where its 'valid' to call blocking primitives (and mutex_lock in particular) and not have TASK_RUNNING, typically such cases are right before we set TASK_RUNNING anyhow. Robustify the code by not assuming this; this has the beneficial side effect of allowing optional code emission for fixing the above might_sleep() false positives. Signed-off-by:
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Wanpeng Li authored
Use nr_cpus_allowed to bail from select_task_rq() when only one cpu can be used, and saves some cycles for pinned tasks. Signed-off-by:
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Wanpeng Li authored
There is no need to do balance during fork since SCHED_DEADLINE tasks can't fork. This patch avoid the SD_BALANCE_FORK check. Signed-off-by:
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Juri Lelli authored
How we deal with updates to exclusive cpusets is currently broken. As an example, suppose we have an exclusive cpuset composed of two cpus: A[cpu0,cpu1]. We can assign SCHED_DEADLINE task to it up to the allowed bandwidth. If we want now to modify cpusetA's cpumask, we have to check that removing a cpu's amount of bandwidth doesn't break AC guarantees. This thing isn't checked in the current code. This patch fixes the problem above, denying an update if the new cpumask won't have enough bandwidth for SCHED_DEADLINE tasks that are currently active. Signed-off-by:
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Juri Lelli authored
Exclusive cpusets are the only way users can restrict SCHED_DEADLINE tasks affinity (performing what is commonly called clustered scheduling). Unfortunately, such thing is currently broken for two reasons: - No check is performed when the user tries to attach a task to an exlusive cpuset (recall that exclusive cpusets have an associated maximum allowed bandwidth). - Bandwidths of source and destination cpusets are not correctly updated after a task is migrated between them. This patch fixes both things at once, as they are opposite faces of the same coin. The check is performed in cpuset_can_attach(), as there aren't any points of failure after that function. The updated is split in two halves. We first reserve bandwidth in the destination cpuset, after we pass the check in cpuset_can_attach(). And we then release bandwidth from the source cpuset when the task's affinity is actually changed. Even if there can be time windows when sched_setattr() may erroneously fail in the source cpuset, we are fine with it, as we can't perfom an atomic update of both cpusets at once. Reported-by:
Daniel Wagner <daniel.wagner@bmw-carit.de> Reported-by:
Vincent Legout <vincent@legout.info> Signed-off-by:
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Wanpeng Li authored
As Kirill mentioned (https://lkml.org/lkml/2013/1/29/118): | If rq has already had 2 or more pushable tasks and we try to add a | pinned task then call of push_rt_task will just waste a time. Just switched pinned task is not able to be pushed. If the rq has had several dl tasks before they have already been considered as candidates to be pushed (or pulled). This patch implements the same behavior as rt class which introduced by commit 10447917 ("sched/rt: Do not try to push tasks if pinned task switches to RT"). Suggested-by:
Kirill V Tkhai <tkhai@yandex.ru> Acked-by:
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Oleg Nesterov authored
task_preempt_count() is pointless if preemption counter is per-cpu, currently this is x86 only. It is only valid if the task is not running, and even in this case the only info it can provide is the state of PREEMPT_ACTIVE bit. Change its single caller to check p->on_rq instead, this should be the same if p->state != TASK_RUNNING, and kill this helper. Signed-off-by:
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Oleg Nesterov authored
Both callers of finish_task_switch() need to recalculate this_rq() and pass it as an argument, plus __schedule() does this again after context_switch(). It would be simpler to call this_rq() once in finish_task_switch() and return the this rq to the callers. Note: probably "int cpu" in __schedule() should die; it is not used and both rcu_note_context_switch() and wq_worker_sleeping() do not really need this argument. Signed-off-by:
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Oleg Nesterov authored
finish_task_switch() enables preemption, so post_schedule(rq) can be called on the wrong (and even dead) CPU. Afaics, nothing really bad can happen, but in this case we can wrongly clear rq->post_schedule on that CPU. And this simply looks wrong in any case. Signed-off-by:
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Oleg Nesterov authored
task_preempt_count() has nothing to do with the actual preempt counter, thread_info->saved_preempt_count is only valid right after switch_to(). __trace_sched_switch_state() can use preempt_count(), prev is still the current task when trace_sched_switch() is called. Signed-off-by:
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Rik van Riel authored
In pseudo-interleaved numa_groups, all tasks try to relocate to the group's preferred_nid. When a group is spread across multiple NUMA nodes, this can lead to tasks swapping their location with other tasks inside the same group, instead of swapping location with tasks from other NUMA groups. This can keep NUMA groups from converging. Examining all nodes, when dealing with a task in a pseudo-interleaved NUMA group, avoids this problem. Note that only CPUs in nodes that improve the task or group score are examined, so the loop isn't too bad. Tested-by:
Vinod Chegu <chegu_vinod@hp.com> Signed-off-by:
Rik van Riel <riel@redhat.com> Signed-off-by:
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Rik van Riel authored
On systems with complex NUMA topologies, the node scoring is adjusted to allow workloads to converge on nodes that are near each other. The way a task group's preferred nid is determined needs to be adjusted, in order for the preferred_nid to be consistent with group_weight scoring. This ensures that we actually try to converge workloads on adjacent nodes. Signed-off-by:
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Rik van Riel authored
In order to do task placement on systems with complex NUMA topologies, it is necessary to count the faults on nodes nearby the node that is being examined for a potential move. In case of a system with a backplane interconnect, we are dealing with groups of NUMA nodes; each of the nodes within a group is the same number of hops away from nodes in other groups in the system. Optimal placement on this topology is achieved by counting all nearby nodes equally. When comparing nodes A and B at distance N, nearby nodes are those at distances smaller than N from nodes A or B. Placement strategy on a system with a glueless mesh NUMA topology needs to be different, because there are no natural groups of nodes determined by the hardware. Instead, when dealing with two nodes A and B at distance N, N >= 2, there will be intermediate nodes at distance < N from both nodes A and B. Good placement can be achieved by right shifting the faults on nearby nodes by the number of hops from the node being scored. In this context, a nearby node is any node less than the maximum distance in the system away from the node. Those nodes are skipped for efficiency reasons, there is no real policy reason to do so. Placement policy on directly connected NUMA systems is not affected. Signed-off-by:
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Rik van Riel authored
Preparatory patch for adding NUMA placement on systems with complex NUMA topology. Also fix a potential divide by zero in group_weight() Signed-off-by:
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Rik van Riel authored
Smaller NUMA systems tend to have all NUMA nodes directly connected to each other. This includes the degenerate case of a system with just one node, ie. a non-NUMA system. Larger systems can have two kinds of NUMA topology, which affects how tasks and memory should be placed on the system. On glueless mesh systems, nodes that are not directly connected to each other will bounce traffic through intermediary nodes. Task groups can be run closer to each other by moving tasks from a node to an intermediary node between it and the task's preferred node. On NUMA systems with backplane controllers, the intermediary hops are incapable of running programs. This creates "islands" of nodes that are at an equal distance to anywhere else in the system. Each kind of topology requires a slightly different placement algorithm; this patch provides the mechanism to detect the kind of NUMA topology of a system. Signed-off-by:
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Rik van Riel authored
Export some information that is necessary to do placement of tasks on systems with multi-level NUMA topologies. Signed-off-by:
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