A crash happened while I was playing with deadline PI rtmutex.
BUG: unable to handle kernel NULL pointer dereference at 0000000000000018
IP: [<ffffffff810eeb8f>] rt_mutex_get_top_task+0x1f/0x30
PGD 232a75067 PUD 230947067 PMD 0
Oops: 0000 [#1] SMP
CPU: 1 PID: 10994 Comm: a.out Not tainted
Call Trace:
[<ffffffff810b658c>] enqueue_task+0x2c/0x80
[<ffffffff810ba763>] activate_task+0x23/0x30
[<ffffffff810d0ab5>] pull_dl_task+0x1d5/0x260
[<ffffffff810d0be6>] pre_schedule_dl+0x16/0x20
[<ffffffff8164e783>] __schedule+0xd3/0x900
[<ffffffff8164efd9>] schedule+0x29/0x70
[<ffffffff8165035b>] __rt_mutex_slowlock+0x4b/0xc0
[<ffffffff81650501>] rt_mutex_slowlock+0xd1/0x190
[<ffffffff810eeb33>] rt_mutex_timed_lock+0x53/0x60
[<ffffffff810ecbfc>] futex_lock_pi.isra.18+0x28c/0x390
[<ffffffff810ed8b0>] do_futex+0x190/0x5b0
[<ffffffff810edd50>] SyS_futex+0x80/0x180
This is because rt_mutex_enqueue_pi() and rt_mutex_dequeue_pi()
are only protected by pi_lock when operating pi waiters, while
rt_mutex_get_top_task(), will access them with rq lock held but
not holding pi_lock.
In order to tackle it, we introduce new "pi_top_task" pointer
cached in task_struct, and add new rt_mutex_update_top_task()
to update its value, it can be called by rt_mutex_setprio()
which held both owner's pi_lock and rq lock. Thus "pi_top_task"
can be safely accessed by enqueue_task_dl() under rq lock.
Originally-From: Peter Zijlstra <peterz@infradead.org>
Signed-off-by: Xunlei Pang <xlpang@redhat.com>
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Acked-by: Steven Rostedt <rostedt@goodmis.org>
Reviewed-by: Thomas Gleixner <tglx@linutronix.de>
Cc: juri.lelli@arm.com
Cc: bigeasy@linutronix.de
Cc: mathieu.desnoyers@efficios.com
Cc: jdesfossez@efficios.com
Cc: bristot@redhat.com
Link: http://lkml.kernel.org/r/20170323150216.157682758@infradead.org
Signed-off-by: Thomas Gleixner <tglx@linutronix.de>
We should deboost before waking the high-priority task, such that we
don't run two tasks with the same "state" (priority, deadline,
sched_class, etc).
In order to make sure the boosting task doesn't start running between
unlock and deboost (due to 'spurious' wakeup), we move the deboost
under the wait_lock, that way its serialized against the wait loop in
__rt_mutex_slowlock().
Doing the deboost early can however lead to priority-inversion if
current would get preempted after the deboost but before waking our
high-prio task, hence we disable preemption before doing deboost, and
enabling it after the wake up is over.
This gets us the right semantic order, but most importantly however;
this change ensures pointer stability for the next patch, where we
have rt_mutex_setprio() cache a pointer to the top-most waiter task.
If we, as before this change, do the wakeup first and then deboost,
this pointer might point into thin air.
[peterz: Changelog + patch munging]
Suggested-by: Peter Zijlstra <peterz@infradead.org>
Signed-off-by: Xunlei Pang <xlpang@redhat.com>
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Acked-by: Steven Rostedt <rostedt@goodmis.org>
Cc: juri.lelli@arm.com
Cc: bigeasy@linutronix.de
Cc: mathieu.desnoyers@efficios.com
Cc: jdesfossez@efficios.com
Cc: bristot@redhat.com
Link: http://lkml.kernel.org/r/20170323150216.110065320@infradead.org
Signed-off-by: Thomas Gleixner <tglx@linutronix.de>
The main PELT function ___update_load_avg(), which implements the
accumulation and progression of the geometric average series, is
implemented along the following lines for the scenario where the time
delta spans all 3 possible sections (see figure below):
1. add the remainder of the last incomplete period
2. decay old sum
3. accumulate new sum in full periods since last_update_time
4. accumulate the current incomplete period
5. update averages
Or:
d1 d2 d3
^ ^ ^
| | |
|<->|<----------------->|<--->|
... |---x---|------| ... |------|-----x (now)
load_sum' = (load_sum + weight * scale * d1) * y^(p+1) + (1,2)
p
weight * scale * 1024 * \Sum y^n + (3)
n=1
weight * scale * d3 * y^0 (4)
load_avg' = load_sum' / LOAD_AVG_MAX (5)
Where:
d1 - is the delta part completing the remainder of the last
incomplete period,
d2 - is the delta part spannind complete periods, and
d3 - is the delta part starting the current incomplete period.
We can simplify the code in two steps; the first step is to separate
the first term into new and old parts like:
(load_sum + weight * scale * d1) * y^(p+1) = load_sum * y^(p+1) +
weight * scale * d1 * y^(p+1)
Once we've done that, its easy to see that all new terms carry the
common factors:
weight * scale
If we factor those out, we arrive at the form:
load_sum' = load_sum * y^(p+1) +
weight * scale * (d1 * y^(p+1) +
p
1024 * \Sum y^n +
n=1
d3 * y^0)
Which results in a simpler, smaller and faster implementation.
Signed-off-by: Yuyang Du <yuyang.du@intel.com>
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: bsegall@google.com
Cc: dietmar.eggemann@arm.com
Cc: matt@codeblueprint.co.uk
Cc: morten.rasmussen@arm.com
Cc: pjt@google.com
Cc: umgwanakikbuti@gmail.com
Cc: vincent.guittot@linaro.org
Link: http://lkml.kernel.org/r/1486935863-25251-3-git-send-email-yuyang.du@intel.com
Signed-off-by: Ingo Molnar <mingo@kernel.org>
The __update_load_avg() function is an __always_inline because its
used with constant propagation to generate different variants of the
code without having to duplicate it (which would be prone to bugs).
Explicitly instantiate the 3 variants.
Note that most of this is called from rather hot paths, so reducing
branches is good.
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: linux-kernel@vger.kernel.org
Signed-off-by: Ingo Molnar <mingo@kernel.org>
Dmitry noted that the new atomic_try_cmpxchg() primitive is broken when
the old pointer doesn't point to the local stack.
He writes:
"Consider a classical lock-free stack push:
node->next = atomic_read(&head);
do {
} while (!atomic_try_cmpxchg(&head, &node->next, node));
This code is broken with the current implementation, the problem is
with unconditional update of *__po.
In case of success it writes the same value back into *__po, but in
case of cmpxchg success we might have lose ownership of some memory
locations and potentially over what __po has pointed to. The same
holds for the re-read of *__po. "
He also points out that this makes it surprisingly different from the
similar C/C++ atomic operation.
After investigating the code-gen differences caused by this patch; and
a number of alternatives (Linus dislikes this interface lots), we
arrived at these results (size x86_64-defconfig/vmlinux):
GCC-6.3.0:
10735757 cmpxchg
10726413 try_cmpxchg
10730509 try_cmpxchg + patch
10730445 try_cmpxchg-linus
GCC-7 (20170327):
10709514 cmpxchg
10704266 try_cmpxchg
10704266 try_cmpxchg + patch
10704394 try_cmpxchg-linus
From this we see that the patch has the advantage of better code-gen
on GCC-7 and keeps the interface roughly consistent with the C
language variant.
Reported-by: Dmitry Vyukov <dvyukov@google.com>
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: Andrew Morton <akpm@linux-foundation.org>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: linux-kernel@vger.kernel.org
Fixes: a9ebf306f5 ("locking/atomic: Introduce atomic_try_cmpxchg()")
Signed-off-by: Ingo Molnar <mingo@kernel.org>
If the child domain prefers tasks to go siblings, the local group could
end up pulling tasks to itself even if the local group is almost equally
loaded as the source group.
Lets assume a 4 core,smt==2 machine running 5 thread ebizzy workload.
Everytime, local group has capacity and source group has atleast 2 threads,
local group tries to pull the task. This causes the threads to constantly
move between different cores. This is even more profound if the cores have
more threads, like in Power 8, smt 8 mode.
Fix this by only allowing local group to pull a task, if the source group
has more number of tasks than the local group.
Here are the relevant perf stat numbers of a 22 core,smt 8 Power 8 machine.
Without patch:
Performance counter stats for 'ebizzy -t 22 -S 100' (5 runs):
1,440 context-switches # 0.001 K/sec ( +- 1.26% )
366 cpu-migrations # 0.000 K/sec ( +- 5.58% )
3,933 page-faults # 0.002 K/sec ( +- 11.08% )
Performance counter stats for 'ebizzy -t 48 -S 100' (5 runs):
6,287 context-switches # 0.001 K/sec ( +- 3.65% )
3,776 cpu-migrations # 0.001 K/sec ( +- 4.84% )
5,702 page-faults # 0.001 K/sec ( +- 9.36% )
Performance counter stats for 'ebizzy -t 96 -S 100' (5 runs):
8,776 context-switches # 0.001 K/sec ( +- 0.73% )
2,790 cpu-migrations # 0.000 K/sec ( +- 0.98% )
10,540 page-faults # 0.001 K/sec ( +- 3.12% )
With patch:
Performance counter stats for 'ebizzy -t 22 -S 100' (5 runs):
1,133 context-switches # 0.001 K/sec ( +- 4.72% )
123 cpu-migrations # 0.000 K/sec ( +- 3.42% )
3,858 page-faults # 0.002 K/sec ( +- 8.52% )
Performance counter stats for 'ebizzy -t 48 -S 100' (5 runs):
2,169 context-switches # 0.000 K/sec ( +- 6.19% )
189 cpu-migrations # 0.000 K/sec ( +- 12.75% )
5,917 page-faults # 0.001 K/sec ( +- 8.09% )
Performance counter stats for 'ebizzy -t 96 -S 100' (5 runs):
5,333 context-switches # 0.001 K/sec ( +- 5.91% )
506 cpu-migrations # 0.000 K/sec ( +- 3.35% )
10,792 page-faults # 0.001 K/sec ( +- 7.75% )
Which show that in these workloads CPU migrations get reduced significantly.
Signed-off-by: Srikar Dronamraju <srikar@linux.vnet.ibm.com>
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: Vincent Guittot <vincent.guittot@linaro.org>
Link: http://lkml.kernel.org/r/1490205470-10249-1-git-send-email-srikar@linux.vnet.ibm.com
Signed-off-by: Ingo Molnar <mingo@kernel.org>
Since commit 383776fa75 ("locking/lockdep: Handle statically initialized
PER_CPU locks properly") we try to collapse per-cpu locks into a single
class by giving them all the same key. For this key we choose the canonical
address of the per-cpu object, which would be the offset into the per-cpu
area.
This has two problems:
- there is a case where we run !0 lock->key through static_obj() and
expect this to pass; it doesn't for canonical pointers.
- 0 is a valid canonical address.
Cure both issues by redefining the canonical address as the address of the
per-cpu variable on the boot CPU.
Since I didn't want to rely on CPU0 being the boot-cpu, or even existing at
all, track the boot CPU in a variable.
Fixes: 383776fa75 ("locking/lockdep: Handle statically initialized PER_CPU locks properly")
Reported-by: kernel test robot <fengguang.wu@intel.com>
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Tested-by: Borislav Petkov <bp@suse.de>
Cc: Sebastian Andrzej Siewior <bigeasy@linutronix.de>
Cc: linux-mm@kvack.org
Cc: wfg@linux.intel.com
Cc: kernel test robot <fengguang.wu@intel.com>
Cc: LKP <lkp@01.org>
Link: http://lkml.kernel.org/r/20170320114108.kbvcsuepem45j5cr@hirez.programming.kicks-ass.net
Signed-off-by: Thomas Gleixner <tglx@linutronix.de>
When PREEMPT_RT_FULL does the spinlock -> rt_mutex substitution the PI
chain code will (falsely) report a deadlock and BUG.
The problem is that it hold hb->lock (now an rt_mutex) while doing
task_blocks_on_rt_mutex on the futex's pi_state::rtmutex. This, when
interleaved just right with futex_unlock_pi() leads it to believe to see an
AB-BA deadlock.
Task1 (holds rt_mutex, Task2 (does FUTEX_LOCK_PI)
does FUTEX_UNLOCK_PI)
lock hb->lock
lock rt_mutex (as per start_proxy)
lock hb->lock
Which is a trivial AB-BA.
It is not an actual deadlock, because it won't be holding hb->lock by the
time it actually blocks on the rt_mutex, but the chainwalk code doesn't
know that and it would be a nightmare to handle this gracefully.
To avoid this problem, do the same as in futex_unlock_pi() and drop
hb->lock after acquiring wait_lock. This still fully serializes against
futex_unlock_pi(), since adding to the wait_list does the very same lock
dance, and removing it holds both locks.
Aside of solving the RT problem this makes the lock and unlock mechanism
symetric and reduces the hb->lock held time.
Reported-and-tested-by: Sebastian Andrzej Siewior <bigeasy@linutronix.de>
Suggested-by: Thomas Gleixner <tglx@linutronix.de>
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: juri.lelli@arm.com
Cc: xlpang@redhat.com
Cc: rostedt@goodmis.org
Cc: mathieu.desnoyers@efficios.com
Cc: jdesfossez@efficios.com
Cc: dvhart@infradead.org
Cc: bristot@redhat.com
Link: http://lkml.kernel.org/r/20170322104152.161341537@infradead.org
Signed-off-by: Thomas Gleixner <tglx@linutronix.de>
By changing futex_lock_pi() to use rt_mutex_*_proxy_lock() all wait_list
modifications are done under both hb->lock and wait_lock.
This closes the obvious interleave pattern between futex_lock_pi() and
futex_unlock_pi(), but not entirely so. See below:
Before:
futex_lock_pi() futex_unlock_pi()
unlock hb->lock
lock hb->lock
unlock hb->lock
lock rt_mutex->wait_lock
unlock rt_mutex_wait_lock
-EAGAIN
lock rt_mutex->wait_lock
list_add
unlock rt_mutex->wait_lock
schedule()
lock rt_mutex->wait_lock
list_del
unlock rt_mutex->wait_lock
<idem>
-EAGAIN
lock hb->lock
After:
futex_lock_pi() futex_unlock_pi()
lock hb->lock
lock rt_mutex->wait_lock
list_add
unlock rt_mutex->wait_lock
unlock hb->lock
schedule()
lock hb->lock
unlock hb->lock
lock hb->lock
lock rt_mutex->wait_lock
list_del
unlock rt_mutex->wait_lock
lock rt_mutex->wait_lock
unlock rt_mutex_wait_lock
-EAGAIN
unlock hb->lock
It does however solve the earlier starvation/live-lock scenario which got
introduced with the -EAGAIN since unlike the before scenario; where the
-EAGAIN happens while futex_unlock_pi() doesn't hold any locks; in the
after scenario it happens while futex_unlock_pi() actually holds a lock,
and then it is serialized on that lock.
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: juri.lelli@arm.com
Cc: bigeasy@linutronix.de
Cc: xlpang@redhat.com
Cc: rostedt@goodmis.org
Cc: mathieu.desnoyers@efficios.com
Cc: jdesfossez@efficios.com
Cc: dvhart@infradead.org
Cc: bristot@redhat.com
Link: http://lkml.kernel.org/r/20170322104152.062785528@infradead.org
Signed-off-by: Thomas Gleixner <tglx@linutronix.de>
With the ultimate goal of keeping rt_mutex wait_list and futex_q waiters
consistent it's necessary to split 'rt_mutex_futex_lock()' into finer
parts, such that only the actual blocking can be done without hb->lock
held.
Split split_mutex_finish_proxy_lock() into two parts, one that does the
blocking and one that does remove_waiter() when the lock acquire failed.
When the rtmutex was acquired successfully the waiter can be removed in the
acquisiton path safely, since there is no concurrency on the lock owner.
This means that, except for futex_lock_pi(), all wait_list modifications
are done with both hb->lock and wait_lock held.
[bigeasy@linutronix.de: fix for futex_requeue_pi_signal_restart]
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: juri.lelli@arm.com
Cc: bigeasy@linutronix.de
Cc: xlpang@redhat.com
Cc: rostedt@goodmis.org
Cc: mathieu.desnoyers@efficios.com
Cc: jdesfossez@efficios.com
Cc: dvhart@infradead.org
Cc: bristot@redhat.com
Link: http://lkml.kernel.org/r/20170322104152.001659630@infradead.org
Signed-off-by: Thomas Gleixner <tglx@linutronix.de>
There is a weird state in the futex_unlock_pi() path when it interleaves
with a concurrent futex_lock_pi() at the point where it drops hb->lock.
In this case, it can happen that the rt_mutex wait_list and the futex_q
disagree on pending waiters, in particular rt_mutex will find no pending
waiters where futex_q thinks there are. In this case the rt_mutex unlock
code cannot assign an owner.
The futex side fixup code has to cleanup the inconsistencies with quite a
bunch of interesting corner cases.
Simplify all this by changing wake_futex_pi() to return -EAGAIN when this
situation occurs. This then gives the futex_lock_pi() code the opportunity
to continue and the retried futex_unlock_pi() will now observe a coherent
state.
The only problem is that this breaks RT timeliness guarantees. That
is, consider the following scenario:
T1 and T2 are both pinned to CPU0. prio(T2) > prio(T1)
CPU0
T1
lock_pi()
queue_me() <- Waiter is visible
preemption
T2
unlock_pi()
loops with -EAGAIN forever
Which is undesirable for PI primitives. Future patches will rectify
this.
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: juri.lelli@arm.com
Cc: bigeasy@linutronix.de
Cc: xlpang@redhat.com
Cc: rostedt@goodmis.org
Cc: mathieu.desnoyers@efficios.com
Cc: jdesfossez@efficios.com
Cc: dvhart@infradead.org
Cc: bristot@redhat.com
Link: http://lkml.kernel.org/r/20170322104151.850383690@infradead.org
Signed-off-by: Thomas Gleixner <tglx@linutronix.de>
Add a new cmpxchg interface:
bool try_cmpxchg(u{8,16,32,64} *ptr, u{8,16,32,64} *val, u{8,16,32,64} new);
Where the boolean returns the result of the compare; and thus if the
exchange happened; and in case of failure, the new value of *ptr is
returned in *val.
This allows simplification/improvement of loops like:
for (;;) {
new = val $op $imm;
old = cmpxchg(ptr, val, new);
if (old == val)
break;
val = old;
}
into:
do {
} while (!try_cmpxchg(ptr, &val, val $op $imm));
while also generating better code (GCC6 and onwards).
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: Andrew Morton <akpm@linux-foundation.org>
Cc: Andy Lutomirski <luto@kernel.org>
Cc: Borislav Petkov <bp@alien8.de>
Cc: Brian Gerst <brgerst@gmail.com>
Cc: Denys Vlasenko <dvlasenk@redhat.com>
Cc: H. Peter Anvin <hpa@zytor.com>
Cc: Josh Poimboeuf <jpoimboe@redhat.com>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: linux-kernel@vger.kernel.org
Signed-off-by: Ingo Molnar <mingo@kernel.org>
A regression of the FTQ noise has been reported by Ying Huang,
on the following hardware:
8 threads Intel(R) Core(TM)i7-4770 CPU @ 3.40GHz with 8G memory
... which was caused by this commit:
commit 4e5160766f ("sched/fair: Propagate asynchrous detach")
The only part of the patch that can increase the noise is the update
of blocked load of group entity in update_blocked_averages().
We can optimize this call and skip the update of group entity if its load
and utilization are already null and there is no pending propagation of load
in the task group.
This optimization partly restores the noise score. A more agressive
optimization has been tried but has shown worse score.
Reported-by: ying.huang@linux.intel.com
Signed-off-by: Vincent Guittot <vincent.guittot@linaro.org>
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: dietmar.eggemann@arm.com
Cc: ying.huang@intel.com
Fixes: 4e5160766f ("sched/fair: Propagate asynchrous detach")
Link: http://lkml.kernel.org/r/1489758442-2877-1-git-send-email-vincent.guittot@linaro.org
[ Fixed typos, improved layout. ]
Signed-off-by: Ingo Molnar <mingo@kernel.org>
If a PER_CPU struct which contains a spin_lock is statically initialized
via:
DEFINE_PER_CPU(struct foo, bla) = {
.lock = __SPIN_LOCK_UNLOCKED(bla.lock)
};
then lockdep assigns a seperate key to each lock because the logic for
assigning a key to statically initialized locks is to use the address as
the key. With per CPU locks the address is obvioulsy different on each CPU.
That's wrong, because all locks should have the same key.
To solve this the following modifications are required:
1) Extend the is_kernel/module_percpu_addr() functions to hand back the
canonical address of the per CPU address, i.e. the per CPU address
minus the per CPU offset.
2) Check the lock address with these functions and if the per CPU check
matches use the returned canonical address as the lock key, so all per
CPU locks have the same key.
3) Move the static_obj(key) check into look_up_lock_class() so this check
can be avoided for statically initialized per CPU locks. That's
required because the canonical address fails the static_obj(key) check
for obvious reasons.
Reported-by: Mike Galbraith <efault@gmx.de>
Signed-off-by: Thomas Gleixner <tglx@linutronix.de>
[ Merged Dan's fixups for !MODULES and !SMP into this patch. ]
Signed-off-by: Sebastian Andrzej Siewior <bigeasy@linutronix.de>
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: Andrew Morton <akpm@linux-foundation.org>
Cc: Dan Murphy <dmurphy@ti.com>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Paul E. McKenney <paulmck@linux.vnet.ibm.com>
Cc: Peter Zijlstra <peterz@infradead.org>
Link: http://lkml.kernel.org/r/20170227143736.pectaimkjkan5kow@linutronix.de
Signed-off-by: Ingo Molnar <mingo@kernel.org>
Instead of relying on deactivate_task() to call update_rq_clock() and
handling the case where it didn't happen (task_on_rq_queued),
unconditionally do update_rq_clock() and skip any further updates.
This also avoids a double update on deactivate_task() + ttwu_local().
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Signed-off-by: Ingo Molnar <mingo@kernel.org>
Currently {en,de}queue_task() do an unconditional update_rq_clock().
However since we want to avoid duplicate updates, so that each
rq->lock section appears atomic in time, we need to be able to skip
these clock updates.
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Signed-off-by: Ingo Molnar <mingo@kernel.org>
The missing update_rq_clock() check can work with partial rq->lock
wrappery, since a missing wrapper can cause the warning to not be
emitted when it should have, but cannot cause the warning to trigger
when it should not have.
The duplicate update_rq_clock() check however can cause false warnings
to trigger. Therefore add more comprehensive rq->lock wrappery.
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Signed-off-by: Ingo Molnar <mingo@kernel.org>
I was testing Daniel's changes with his test case, and tweaked it a
little. Instead of having the runtime equal to the deadline, I
increased the deadline ten fold.
Daniel's test case had:
attr.sched_runtime = 2 * 1000 * 1000; /* 2 ms */
attr.sched_deadline = 2 * 1000 * 1000; /* 2 ms */
attr.sched_period = 2 * 1000 * 1000 * 1000; /* 2 s */
To make it more interesting, I changed it to:
attr.sched_runtime = 2 * 1000 * 1000; /* 2 ms */
attr.sched_deadline = 20 * 1000 * 1000; /* 20 ms */
attr.sched_period = 2 * 1000 * 1000 * 1000; /* 2 s */
The results were rather surprising. The behavior that Daniel's patch
was fixing came back. The task started using much more than .1% of the
CPU. More like 20%.
Looking into this I found that it was due to the dl_entity_overflow()
constantly returning true. That's because it uses the relative period
against relative runtime vs the absolute deadline against absolute
runtime.
runtime / (deadline - t) > dl_runtime / dl_period
There's even a comment mentioning this, and saying that when relative
deadline equals relative period, that the equation is the same as using
deadline instead of period. That comment is backwards! What we really
want is:
runtime / (deadline - t) > dl_runtime / dl_deadline
We care about if the runtime can make its deadline, not its period. And
then we can say "when the deadline equals the period, the equation is
the same as using dl_period instead of dl_deadline".
After correcting this, now when the task gets enqueued, it can throttle
correctly, and Daniel's fix to the throttling of sleeping deadline
tasks works even when the runtime and deadline are not the same.
Signed-off-by: Steven Rostedt (VMware) <rostedt@goodmis.org>
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Reviewed-by: Daniel Bristot de Oliveira <bristot@redhat.com>
Cc: Juri Lelli <juri.lelli@arm.com>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Luca Abeni <luca.abeni@santannapisa.it>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Romulo Silva de Oliveira <romulo.deoliveira@ufsc.br>
Cc: Steven Rostedt <rostedt@goodmis.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: Tommaso Cucinotta <tommaso.cucinotta@sssup.it>
Link: http://lkml.kernel.org/r/02135a27f1ae3fe5fd032568a5a2f370e190e8d7.1488392936.git.bristot@redhat.com
Signed-off-by: Ingo Molnar <mingo@kernel.org>
During the activation, CBS checks if it can reuse the current task's
runtime and period. If the deadline of the task is in the past, CBS
cannot use the runtime, and so it replenishes the task. This rule
works fine for implicit deadline tasks (deadline == period), and the
CBS was designed for implicit deadline tasks. However, a task with
constrained deadline (deadine < period) might be awakened after the
deadline, but before the next period. In this case, replenishing the
task would allow it to run for runtime / deadline. As in this case
deadline < period, CBS enables a task to run for more than the
runtime / period. In a very loaded system, this can cause a domino
effect, making other tasks miss their deadlines.
To avoid this problem, in the activation of a constrained deadline
task after the deadline but before the next period, throttle the
task and set the replenishing timer to the begin of the next period,
unless it is boosted.
Reproducer:
--------------- %< ---------------
int main (int argc, char **argv)
{
int ret;
int flags = 0;
unsigned long l = 0;
struct timespec ts;
struct sched_attr attr;
memset(&attr, 0, sizeof(attr));
attr.size = sizeof(attr);
attr.sched_policy = SCHED_DEADLINE;
attr.sched_runtime = 2 * 1000 * 1000; /* 2 ms */
attr.sched_deadline = 2 * 1000 * 1000; /* 2 ms */
attr.sched_period = 2 * 1000 * 1000 * 1000; /* 2 s */
ts.tv_sec = 0;
ts.tv_nsec = 2000 * 1000; /* 2 ms */
ret = sched_setattr(0, &attr, flags);
if (ret < 0) {
perror("sched_setattr");
exit(-1);
}
for(;;) {
/* XXX: you may need to adjust the loop */
for (l = 0; l < 150000; l++);
/*
* The ideia is to go to sleep right before the deadline
* and then wake up before the next period to receive
* a new replenishment.
*/
nanosleep(&ts, NULL);
}
exit(0);
}
--------------- >% ---------------
On my box, this reproducer uses almost 50% of the CPU time, which is
obviously wrong for a task with 2/2000 reservation.
Signed-off-by: Daniel Bristot de Oliveira <bristot@redhat.com>
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Cc: Juri Lelli <juri.lelli@arm.com>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Luca Abeni <luca.abeni@santannapisa.it>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Romulo Silva de Oliveira <romulo.deoliveira@ufsc.br>
Cc: Steven Rostedt <rostedt@goodmis.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: Tommaso Cucinotta <tommaso.cucinotta@sssup.it>
Link: http://lkml.kernel.org/r/edf58354e01db46bf42df8d2dd32418833f68c89.1488392936.git.bristot@redhat.com
Signed-off-by: Ingo Molnar <mingo@kernel.org>
Currently, the replenishment timer is set to fire at the deadline
of a task. Although that works for implicit deadline tasks because the
deadline is equals to the begin of the next period, that is not correct
for constrained deadline tasks (deadline < period).
For instance:
f.c:
--------------- %< ---------------
int main (void)
{
for(;;);
}
--------------- >% ---------------
# gcc -o f f.c
# trace-cmd record -e sched:sched_switch \
-e syscalls:sys_exit_sched_setattr \
chrt -d --sched-runtime 490000000 \
--sched-deadline 500000000 \
--sched-period 1000000000 0 ./f
# trace-cmd report | grep "{pid of ./f}"
After setting parameters, the task is replenished and continue running
until being throttled:
f-11295 [003] 13322.113776: sys_exit_sched_setattr: 0x0
The task is throttled after running 492318 ms, as expected:
f-11295 [003] 13322.606094: sched_switch: f:11295 [-1] R ==> watchdog/3:32 [0]
But then, the task is replenished 500719 ms after the first
replenishment:
<idle>-0 [003] 13322.614495: sched_switch: swapper/3:0 [120] R ==> f:11295 [-1]
Running for 490277 ms:
f-11295 [003] 13323.104772: sched_switch: f:11295 [-1] R ==> swapper/3:0 [120]
Hence, in the first period, the task runs 2 * runtime, and that is a bug.
During the first replenishment, the next deadline is set one period away.
So the runtime / period starts to be respected. However, as the second
replenishment took place in the wrong instant, the next replenishment
will also be held in a wrong instant of time. Rather than occurring in
the nth period away from the first activation, it is taking place
in the (nth period - relative deadline).
Signed-off-by: Daniel Bristot de Oliveira <bristot@redhat.com>
Signed-off-by: Peter Zijlstra (Intel) <peterz@infradead.org>
Reviewed-by: Luca Abeni <luca.abeni@santannapisa.it>
Reviewed-by: Steven Rostedt (VMware) <rostedt@goodmis.org>
Reviewed-by: Juri Lelli <juri.lelli@arm.com>
Cc: Linus Torvalds <torvalds@linux-foundation.org>
Cc: Mike Galbraith <efault@gmx.de>
Cc: Peter Zijlstra <peterz@infradead.org>
Cc: Romulo Silva de Oliveira <romulo.deoliveira@ufsc.br>
Cc: Steven Rostedt <rostedt@goodmis.org>
Cc: Thomas Gleixner <tglx@linutronix.de>
Cc: Tommaso Cucinotta <tommaso.cucinotta@sssup.it>
Link: http://lkml.kernel.org/r/ac50d89887c25285b47465638354b63362f8adff.1488392936.git.bristot@redhat.com
Signed-off-by: Ingo Molnar <mingo@kernel.org>
