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jdk-24/hotspot/src/share/vm/runtime/synchronizer.cpp
Xiaobin Lu 62fa852746 6699669: Hotspot server leaves synchronized block with monitor in bad state 
Remove usage of _highest_lock field in Thread so that is_lock_owned won't depend on the correct update of that field.

Reviewed-by: never, dice, acorn
2009-04-06 15:47:39 -07:00

4716 lines
191 KiB
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/*
* Copyright 1998-2009 Sun Microsystems, Inc. All Rights Reserved.
* DO NOT ALTER OR REMOVE COPYRIGHT NOTICES OR THIS FILE HEADER.
*
* This code is free software; you can redistribute it and/or modify it
* under the terms of the GNU General Public License version 2 only, as
* published by the Free Software Foundation.
*
* This code is distributed in the hope that it will be useful, but WITHOUT
* ANY WARRANTY; without even the implied warranty of MERCHANTABILITY or
* FITNESS FOR A PARTICULAR PURPOSE. See the GNU General Public License
* version 2 for more details (a copy is included in the LICENSE file that
* accompanied this code).
*
* You should have received a copy of the GNU General Public License version
* 2 along with this work; if not, write to the Free Software Foundation,
* Inc., 51 Franklin St, Fifth Floor, Boston, MA 02110-1301 USA.
*
* Please contact Sun Microsystems, Inc., 4150 Network Circle, Santa Clara,
* CA 95054 USA or visit www.sun.com if you need additional information or
* have any questions.
*
*/
# include "incls/_precompiled.incl"
# include "incls/_synchronizer.cpp.incl"
#if defined(__GNUC__) && !defined(IA64)
// Need to inhibit inlining for older versions of GCC to avoid build-time failures
#define ATTR __attribute__((noinline))
#else
#define ATTR
#endif
// Native markword accessors for synchronization and hashCode().
//
// The "core" versions of monitor enter and exit reside in this file.
// The interpreter and compilers contain specialized transliterated
// variants of the enter-exit fast-path operations. See i486.ad fast_lock(),
// for instance. If you make changes here, make sure to modify the
// interpreter, and both C1 and C2 fast-path inline locking code emission.
//
// TODO: merge the objectMonitor and synchronizer classes.
//
// -----------------------------------------------------------------------------
#ifdef DTRACE_ENABLED
// Only bother with this argument setup if dtrace is available
// TODO-FIXME: probes should not fire when caller is _blocked. assert() accordingly.
HS_DTRACE_PROBE_DECL5(hotspot, monitor__wait,
jlong, uintptr_t, char*, int, long);
HS_DTRACE_PROBE_DECL4(hotspot, monitor__waited,
jlong, uintptr_t, char*, int);
HS_DTRACE_PROBE_DECL4(hotspot, monitor__notify,
jlong, uintptr_t, char*, int);
HS_DTRACE_PROBE_DECL4(hotspot, monitor__notifyAll,
jlong, uintptr_t, char*, int);
HS_DTRACE_PROBE_DECL4(hotspot, monitor__contended__enter,
jlong, uintptr_t, char*, int);
HS_DTRACE_PROBE_DECL4(hotspot, monitor__contended__entered,
jlong, uintptr_t, char*, int);
HS_DTRACE_PROBE_DECL4(hotspot, monitor__contended__exit,
jlong, uintptr_t, char*, int);
#define DTRACE_MONITOR_PROBE_COMMON(klassOop, thread) \
char* bytes = NULL; \
int len = 0; \
jlong jtid = SharedRuntime::get_java_tid(thread); \
symbolOop klassname = ((oop)(klassOop))->klass()->klass_part()->name(); \
if (klassname != NULL) { \
bytes = (char*)klassname->bytes(); \
len = klassname->utf8_length(); \
}
#define DTRACE_MONITOR_WAIT_PROBE(monitor, klassOop, thread, millis) \
{ \
if (DTraceMonitorProbes) { \
DTRACE_MONITOR_PROBE_COMMON(klassOop, thread); \
HS_DTRACE_PROBE5(hotspot, monitor__wait, jtid, \
(monitor), bytes, len, (millis)); \
} \
}
#define DTRACE_MONITOR_PROBE(probe, monitor, klassOop, thread) \
{ \
if (DTraceMonitorProbes) { \
DTRACE_MONITOR_PROBE_COMMON(klassOop, thread); \
HS_DTRACE_PROBE4(hotspot, monitor__##probe, jtid, \
(uintptr_t)(monitor), bytes, len); \
} \
}
#else // ndef DTRACE_ENABLED
#define DTRACE_MONITOR_WAIT_PROBE(klassOop, thread, millis, mon) {;}
#define DTRACE_MONITOR_PROBE(probe, klassOop, thread, mon) {;}
#endif // ndef DTRACE_ENABLED
// ObjectWaiter serves as a "proxy" or surrogate thread.
// TODO-FIXME: Eliminate ObjectWaiter and use the thread-specific
// ParkEvent instead. Beware, however, that the JVMTI code
// knows about ObjectWaiters, so we'll have to reconcile that code.
// See next_waiter(), first_waiter(), etc.
class ObjectWaiter : public StackObj {
public:
enum TStates { TS_UNDEF, TS_READY, TS_RUN, TS_WAIT, TS_ENTER, TS_CXQ } ;
enum Sorted { PREPEND, APPEND, SORTED } ;
ObjectWaiter * volatile _next;
ObjectWaiter * volatile _prev;
Thread* _thread;
ParkEvent * _event;
volatile int _notified ;
volatile TStates TState ;
Sorted _Sorted ; // List placement disposition
bool _active ; // Contention monitoring is enabled
public:
ObjectWaiter(Thread* thread) {
_next = NULL;
_prev = NULL;
_notified = 0;
TState = TS_RUN ;
_thread = thread;
_event = thread->_ParkEvent ;
_active = false;
assert (_event != NULL, "invariant") ;
}
void wait_reenter_begin(ObjectMonitor *mon) {
JavaThread *jt = (JavaThread *)this->_thread;
_active = JavaThreadBlockedOnMonitorEnterState::wait_reenter_begin(jt, mon);
}
void wait_reenter_end(ObjectMonitor *mon) {
JavaThread *jt = (JavaThread *)this->_thread;
JavaThreadBlockedOnMonitorEnterState::wait_reenter_end(jt, _active);
}
};
enum ManifestConstants {
ClearResponsibleAtSTW = 0,
MaximumRecheckInterval = 1000
} ;
#undef TEVENT
#define TEVENT(nom) {if (SyncVerbose) FEVENT(nom); }
#define FEVENT(nom) { static volatile int ctr = 0 ; int v = ++ctr ; if ((v & (v-1)) == 0) { ::printf (#nom " : %d \n", v); ::fflush(stdout); }}
#undef TEVENT
#define TEVENT(nom) {;}
// Performance concern:
// OrderAccess::storestore() calls release() which STs 0 into the global volatile
// OrderAccess::Dummy variable. This store is unnecessary for correctness.
// Many threads STing into a common location causes considerable cache migration
// or "sloshing" on large SMP system. As such, I avoid using OrderAccess::storestore()
// until it's repaired. In some cases OrderAccess::fence() -- which incurs local
// latency on the executing processor -- is a better choice as it scales on SMP
// systems. See http://blogs.sun.com/dave/entry/biased_locking_in_hotspot for a
// discussion of coherency costs. Note that all our current reference platforms
// provide strong ST-ST order, so the issue is moot on IA32, x64, and SPARC.
//
// As a general policy we use "volatile" to control compiler-based reordering
// and explicit fences (barriers) to control for architectural reordering performed
// by the CPU(s) or platform.
static int MBFence (int x) { OrderAccess::fence(); return x; }
struct SharedGlobals {
// These are highly shared mostly-read variables.
// To avoid false-sharing they need to be the sole occupants of a $ line.
double padPrefix [8];
volatile int stwRandom ;
volatile int stwCycle ;
// Hot RW variables -- Sequester to avoid false-sharing
double padSuffix [16];
volatile int hcSequence ;
double padFinal [8] ;
} ;
static SharedGlobals GVars ;
// Tunables ...
// The knob* variables are effectively final. Once set they should
// never be modified hence. Consider using __read_mostly with GCC.
static int Knob_LogSpins = 0 ; // enable jvmstat tally for spins
static int Knob_HandOff = 0 ;
static int Knob_Verbose = 0 ;
static int Knob_ReportSettings = 0 ;
static int Knob_SpinLimit = 5000 ; // derived by an external tool -
static int Knob_SpinBase = 0 ; // Floor AKA SpinMin
static int Knob_SpinBackOff = 0 ; // spin-loop backoff
static int Knob_CASPenalty = -1 ; // Penalty for failed CAS
static int Knob_OXPenalty = -1 ; // Penalty for observed _owner change
static int Knob_SpinSetSucc = 1 ; // spinners set the _succ field
static int Knob_SpinEarly = 1 ;
static int Knob_SuccEnabled = 1 ; // futile wake throttling
static int Knob_SuccRestrict = 0 ; // Limit successors + spinners to at-most-one
static int Knob_MaxSpinners = -1 ; // Should be a function of # CPUs
static int Knob_Bonus = 100 ; // spin success bonus
static int Knob_BonusB = 100 ; // spin success bonus
static int Knob_Penalty = 200 ; // spin failure penalty
static int Knob_Poverty = 1000 ;
static int Knob_SpinAfterFutile = 1 ; // Spin after returning from park()
static int Knob_FixedSpin = 0 ;
static int Knob_OState = 3 ; // Spinner checks thread state of _owner
static int Knob_UsePause = 1 ;
static int Knob_ExitPolicy = 0 ;
static int Knob_PreSpin = 10 ; // 20-100 likely better
static int Knob_ResetEvent = 0 ;
static int BackOffMask = 0 ;
static int Knob_FastHSSEC = 0 ;
static int Knob_MoveNotifyee = 2 ; // notify() - disposition of notifyee
static int Knob_QMode = 0 ; // EntryList-cxq policy - queue discipline
static volatile int InitDone = 0 ;
// hashCode() generation :
//
// Possibilities:
// * MD5Digest of {obj,stwRandom}
// * CRC32 of {obj,stwRandom} or any linear-feedback shift register function.
// * A DES- or AES-style SBox[] mechanism
// * One of the Phi-based schemes, such as:
// 2654435761 = 2^32 * Phi (golden ratio)
// HashCodeValue = ((uintptr_t(obj) >> 3) * 2654435761) ^ GVars.stwRandom ;
// * A variation of Marsaglia's shift-xor RNG scheme.
// * (obj ^ stwRandom) is appealing, but can result
// in undesirable regularity in the hashCode values of adjacent objects
// (objects allocated back-to-back, in particular). This could potentially
// result in hashtable collisions and reduced hashtable efficiency.
// There are simple ways to "diffuse" the middle address bits over the
// generated hashCode values:
//
static inline intptr_t get_next_hash(Thread * Self, oop obj) {
intptr_t value = 0 ;
if (hashCode == 0) {
// This form uses an unguarded global Park-Miller RNG,
// so it's possible for two threads to race and generate the same RNG.
// On MP system we'll have lots of RW access to a global, so the
// mechanism induces lots of coherency traffic.
value = os::random() ;
} else
if (hashCode == 1) {
// This variation has the property of being stable (idempotent)
// between STW operations. This can be useful in some of the 1-0
// synchronization schemes.
intptr_t addrBits = intptr_t(obj) >> 3 ;
value = addrBits ^ (addrBits >> 5) ^ GVars.stwRandom ;
} else
if (hashCode == 2) {
value = 1 ; // for sensitivity testing
} else
if (hashCode == 3) {
value = ++GVars.hcSequence ;
} else
if (hashCode == 4) {
value = intptr_t(obj) ;
} else {
// Marsaglia's xor-shift scheme with thread-specific state
// This is probably the best overall implementation -- we'll
// likely make this the default in future releases.
unsigned t = Self->_hashStateX ;
t ^= (t << 11) ;
Self->_hashStateX = Self->_hashStateY ;
Self->_hashStateY = Self->_hashStateZ ;
Self->_hashStateZ = Self->_hashStateW ;
unsigned v = Self->_hashStateW ;
v = (v ^ (v >> 19)) ^ (t ^ (t >> 8)) ;
Self->_hashStateW = v ;
value = v ;
}
value &= markOopDesc::hash_mask;
if (value == 0) value = 0xBAD ;
assert (value != markOopDesc::no_hash, "invariant") ;
TEVENT (hashCode: GENERATE) ;
return value;
}
void BasicLock::print_on(outputStream* st) const {
st->print("monitor");
}
void BasicLock::move_to(oop obj, BasicLock* dest) {
// Check to see if we need to inflate the lock. This is only needed
// if an object is locked using "this" lightweight monitor. In that
// case, the displaced_header() is unlocked, because the
// displaced_header() contains the header for the originally unlocked
// object. However the object could have already been inflated. But it
// does not matter, the inflation will just a no-op. For other cases,
// the displaced header will be either 0x0 or 0x3, which are location
// independent, therefore the BasicLock is free to move.
//
// During OSR we may need to relocate a BasicLock (which contains a
// displaced word) from a location in an interpreter frame to a
// new location in a compiled frame. "this" refers to the source
// basiclock in the interpreter frame. "dest" refers to the destination
// basiclock in the new compiled frame. We *always* inflate in move_to().
// The always-Inflate policy works properly, but in 1.5.0 it can sometimes
// cause performance problems in code that makes heavy use of a small # of
// uncontended locks. (We'd inflate during OSR, and then sync performance
// would subsequently plummet because the thread would be forced thru the slow-path).
// This problem has been made largely moot on IA32 by inlining the inflated fast-path
// operations in Fast_Lock and Fast_Unlock in i486.ad.
//
// Note that there is a way to safely swing the object's markword from
// one stack location to another. This avoids inflation. Obviously,
// we need to ensure that both locations refer to the current thread's stack.
// There are some subtle concurrency issues, however, and since the benefit is
// is small (given the support for inflated fast-path locking in the fast_lock, etc)
// we'll leave that optimization for another time.
if (displaced_header()->is_neutral()) {
ObjectSynchronizer::inflate_helper(obj);
// WARNING: We can not put check here, because the inflation
// will not update the displaced header. Once BasicLock is inflated,
// no one should ever look at its content.
} else {
// Typically the displaced header will be 0 (recursive stack lock) or
// unused_mark. Naively we'd like to assert that the displaced mark
// value is either 0, neutral, or 3. But with the advent of the
// store-before-CAS avoidance in fast_lock/compiler_lock_object
// we can find any flavor mark in the displaced mark.
}
// [RGV] The next line appears to do nothing!
intptr_t dh = (intptr_t) displaced_header();
dest->set_displaced_header(displaced_header());
}
// -----------------------------------------------------------------------------
// standard constructor, allows locking failures
ObjectLocker::ObjectLocker(Handle obj, Thread* thread, bool doLock) {
_dolock = doLock;
_thread = thread;
debug_only(if (StrictSafepointChecks) _thread->check_for_valid_safepoint_state(false);)
_obj = obj;
if (_dolock) {
TEVENT (ObjectLocker) ;
ObjectSynchronizer::fast_enter(_obj, &_lock, false, _thread);
}
}
ObjectLocker::~ObjectLocker() {
if (_dolock) {
ObjectSynchronizer::fast_exit(_obj(), &_lock, _thread);
}
}
// -----------------------------------------------------------------------------
PerfCounter * ObjectSynchronizer::_sync_Inflations = NULL ;
PerfCounter * ObjectSynchronizer::_sync_Deflations = NULL ;
PerfCounter * ObjectSynchronizer::_sync_ContendedLockAttempts = NULL ;
PerfCounter * ObjectSynchronizer::_sync_FutileWakeups = NULL ;
PerfCounter * ObjectSynchronizer::_sync_Parks = NULL ;
PerfCounter * ObjectSynchronizer::_sync_EmptyNotifications = NULL ;
PerfCounter * ObjectSynchronizer::_sync_Notifications = NULL ;
PerfCounter * ObjectSynchronizer::_sync_PrivateA = NULL ;
PerfCounter * ObjectSynchronizer::_sync_PrivateB = NULL ;
PerfCounter * ObjectSynchronizer::_sync_SlowExit = NULL ;
PerfCounter * ObjectSynchronizer::_sync_SlowEnter = NULL ;
PerfCounter * ObjectSynchronizer::_sync_SlowNotify = NULL ;
PerfCounter * ObjectSynchronizer::_sync_SlowNotifyAll = NULL ;
PerfCounter * ObjectSynchronizer::_sync_FailedSpins = NULL ;
PerfCounter * ObjectSynchronizer::_sync_SuccessfulSpins = NULL ;
PerfCounter * ObjectSynchronizer::_sync_MonInCirculation = NULL ;
PerfCounter * ObjectSynchronizer::_sync_MonScavenged = NULL ;
PerfLongVariable * ObjectSynchronizer::_sync_MonExtant = NULL ;
// One-shot global initialization for the sync subsystem.
// We could also defer initialization and initialize on-demand
// the first time we call inflate(). Initialization would
// be protected - like so many things - by the MonitorCache_lock.
void ObjectSynchronizer::Initialize () {
static int InitializationCompleted = 0 ;
assert (InitializationCompleted == 0, "invariant") ;
InitializationCompleted = 1 ;
if (UsePerfData) {
EXCEPTION_MARK ;
#define NEWPERFCOUNTER(n) {n = PerfDataManager::create_counter(SUN_RT, #n, PerfData::U_Events,CHECK); }
#define NEWPERFVARIABLE(n) {n = PerfDataManager::create_variable(SUN_RT, #n, PerfData::U_Events,CHECK); }
NEWPERFCOUNTER(_sync_Inflations) ;
NEWPERFCOUNTER(_sync_Deflations) ;
NEWPERFCOUNTER(_sync_ContendedLockAttempts) ;
NEWPERFCOUNTER(_sync_FutileWakeups) ;
NEWPERFCOUNTER(_sync_Parks) ;
NEWPERFCOUNTER(_sync_EmptyNotifications) ;
NEWPERFCOUNTER(_sync_Notifications) ;
NEWPERFCOUNTER(_sync_SlowEnter) ;
NEWPERFCOUNTER(_sync_SlowExit) ;
NEWPERFCOUNTER(_sync_SlowNotify) ;
NEWPERFCOUNTER(_sync_SlowNotifyAll) ;
NEWPERFCOUNTER(_sync_FailedSpins) ;
NEWPERFCOUNTER(_sync_SuccessfulSpins) ;
NEWPERFCOUNTER(_sync_PrivateA) ;
NEWPERFCOUNTER(_sync_PrivateB) ;
NEWPERFCOUNTER(_sync_MonInCirculation) ;
NEWPERFCOUNTER(_sync_MonScavenged) ;
NEWPERFVARIABLE(_sync_MonExtant) ;
#undef NEWPERFCOUNTER
}
}
// Compile-time asserts
// When possible, it's better to catch errors deterministically at
// compile-time than at runtime. The down-side to using compile-time
// asserts is that error message -- often something about negative array
// indices -- is opaque.
#define CTASSERT(x) { int tag[1-(2*!(x))]; printf ("Tag @" INTPTR_FORMAT "\n", (intptr_t)tag); }
void ObjectMonitor::ctAsserts() {
CTASSERT(offset_of (ObjectMonitor, _header) == 0);
}
static int Adjust (volatile int * adr, int dx) {
int v ;
for (v = *adr ; Atomic::cmpxchg (v + dx, adr, v) != v; v = *adr) ;
return v ;
}
// Ad-hoc mutual exclusion primitives: SpinLock and Mux
//
// We employ SpinLocks _only for low-contention, fixed-length
// short-duration critical sections where we're concerned
// about native mutex_t or HotSpot Mutex:: latency.
// The mux construct provides a spin-then-block mutual exclusion
// mechanism.
//
// Testing has shown that contention on the ListLock guarding gFreeList
// is common. If we implement ListLock as a simple SpinLock it's common
// for the JVM to devolve to yielding with little progress. This is true
// despite the fact that the critical sections protected by ListLock are
// extremely short.
//
// TODO-FIXME: ListLock should be of type SpinLock.
// We should make this a 1st-class type, integrated into the lock
// hierarchy as leaf-locks. Critically, the SpinLock structure
// should have sufficient padding to avoid false-sharing and excessive
// cache-coherency traffic.
typedef volatile int SpinLockT ;
void Thread::SpinAcquire (volatile int * adr, const char * LockName) {
if (Atomic::cmpxchg (1, adr, 0) == 0) {
return ; // normal fast-path return
}
// Slow-path : We've encountered contention -- Spin/Yield/Block strategy.
TEVENT (SpinAcquire - ctx) ;
int ctr = 0 ;
int Yields = 0 ;
for (;;) {
while (*adr != 0) {
++ctr ;
if ((ctr & 0xFFF) == 0 || !os::is_MP()) {
if (Yields > 5) {
// Consider using a simple NakedSleep() instead.
// Then SpinAcquire could be called by non-JVM threads
Thread::current()->_ParkEvent->park(1) ;
} else {
os::NakedYield() ;
++Yields ;
}
} else {
SpinPause() ;
}
}
if (Atomic::cmpxchg (1, adr, 0) == 0) return ;
}
}
void Thread::SpinRelease (volatile int * adr) {
assert (*adr != 0, "invariant") ;
OrderAccess::fence() ; // guarantee at least release consistency.
// Roach-motel semantics.
// It's safe if subsequent LDs and STs float "up" into the critical section,
// but prior LDs and STs within the critical section can't be allowed
// to reorder or float past the ST that releases the lock.
*adr = 0 ;
}
// muxAcquire and muxRelease:
//
// * muxAcquire and muxRelease support a single-word lock-word construct.
// The LSB of the word is set IFF the lock is held.
// The remainder of the word points to the head of a singly-linked list
// of threads blocked on the lock.
//
// * The current implementation of muxAcquire-muxRelease uses its own
// dedicated Thread._MuxEvent instance. If we're interested in
// minimizing the peak number of extant ParkEvent instances then
// we could eliminate _MuxEvent and "borrow" _ParkEvent as long
// as certain invariants were satisfied. Specifically, care would need
// to be taken with regards to consuming unpark() "permits".
// A safe rule of thumb is that a thread would never call muxAcquire()
// if it's enqueued (cxq, EntryList, WaitList, etc) and will subsequently
// park(). Otherwise the _ParkEvent park() operation in muxAcquire() could
// consume an unpark() permit intended for monitorenter, for instance.
// One way around this would be to widen the restricted-range semaphore
// implemented in park(). Another alternative would be to provide
// multiple instances of the PlatformEvent() for each thread. One
// instance would be dedicated to muxAcquire-muxRelease, for instance.
//
// * Usage:
// -- Only as leaf locks
// -- for short-term locking only as muxAcquire does not perform
// thread state transitions.
//
// Alternatives:
// * We could implement muxAcquire and muxRelease with MCS or CLH locks
// but with parking or spin-then-park instead of pure spinning.
// * Use Taura-Oyama-Yonenzawa locks.
// * It's possible to construct a 1-0 lock if we encode the lockword as
// (List,LockByte). Acquire will CAS the full lockword while Release
// will STB 0 into the LockByte. The 1-0 scheme admits stranding, so
// acquiring threads use timers (ParkTimed) to detect and recover from
// the stranding window. Thread/Node structures must be aligned on 256-byte
// boundaries by using placement-new.
// * Augment MCS with advisory back-link fields maintained with CAS().
// Pictorially: LockWord -> T1 <-> T2 <-> T3 <-> ... <-> Tn <-> Owner.
// The validity of the backlinks must be ratified before we trust the value.
// If the backlinks are invalid the exiting thread must back-track through the
// the forward links, which are always trustworthy.
// * Add a successor indication. The LockWord is currently encoded as
// (List, LOCKBIT:1). We could also add a SUCCBIT or an explicit _succ variable
// to provide the usual futile-wakeup optimization.
// See RTStt for details.
// * Consider schedctl.sc_nopreempt to cover the critical section.
//
typedef volatile intptr_t MutexT ; // Mux Lock-word
enum MuxBits { LOCKBIT = 1 } ;
void Thread::muxAcquire (volatile intptr_t * Lock, const char * LockName) {
intptr_t w = Atomic::cmpxchg_ptr (LOCKBIT, Lock, 0) ;
if (w == 0) return ;
if ((w & LOCKBIT) == 0 && Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) {
return ;
}
TEVENT (muxAcquire - Contention) ;
ParkEvent * const Self = Thread::current()->_MuxEvent ;
assert ((intptr_t(Self) & LOCKBIT) == 0, "invariant") ;
for (;;) {
int its = (os::is_MP() ? 100 : 0) + 1 ;
// Optional spin phase: spin-then-park strategy
while (--its >= 0) {
w = *Lock ;
if ((w & LOCKBIT) == 0 && Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) {
return ;
}
}
Self->reset() ;
Self->OnList = intptr_t(Lock) ;
// The following fence() isn't _strictly necessary as the subsequent
// CAS() both serializes execution and ratifies the fetched *Lock value.
OrderAccess::fence();
for (;;) {
w = *Lock ;
if ((w & LOCKBIT) == 0) {
if (Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) {
Self->OnList = 0 ; // hygiene - allows stronger asserts
return ;
}
continue ; // Interference -- *Lock changed -- Just retry
}
assert (w & LOCKBIT, "invariant") ;
Self->ListNext = (ParkEvent *) (w & ~LOCKBIT );
if (Atomic::cmpxchg_ptr (intptr_t(Self)|LOCKBIT, Lock, w) == w) break ;
}
while (Self->OnList != 0) {
Self->park() ;
}
}
}
void Thread::muxAcquireW (volatile intptr_t * Lock, ParkEvent * ev) {
intptr_t w = Atomic::cmpxchg_ptr (LOCKBIT, Lock, 0) ;
if (w == 0) return ;
if ((w & LOCKBIT) == 0 && Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) {
return ;
}
TEVENT (muxAcquire - Contention) ;
ParkEvent * ReleaseAfter = NULL ;
if (ev == NULL) {
ev = ReleaseAfter = ParkEvent::Allocate (NULL) ;
}
assert ((intptr_t(ev) & LOCKBIT) == 0, "invariant") ;
for (;;) {
guarantee (ev->OnList == 0, "invariant") ;
int its = (os::is_MP() ? 100 : 0) + 1 ;
// Optional spin phase: spin-then-park strategy
while (--its >= 0) {
w = *Lock ;
if ((w & LOCKBIT) == 0 && Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) {
if (ReleaseAfter != NULL) {
ParkEvent::Release (ReleaseAfter) ;
}
return ;
}
}
ev->reset() ;
ev->OnList = intptr_t(Lock) ;
// The following fence() isn't _strictly necessary as the subsequent
// CAS() both serializes execution and ratifies the fetched *Lock value.
OrderAccess::fence();
for (;;) {
w = *Lock ;
if ((w & LOCKBIT) == 0) {
if (Atomic::cmpxchg_ptr (w|LOCKBIT, Lock, w) == w) {
ev->OnList = 0 ;
// We call ::Release while holding the outer lock, thus
// artificially lengthening the critical section.
// Consider deferring the ::Release() until the subsequent unlock(),
// after we've dropped the outer lock.
if (ReleaseAfter != NULL) {
ParkEvent::Release (ReleaseAfter) ;
}
return ;
}
continue ; // Interference -- *Lock changed -- Just retry
}
assert (w & LOCKBIT, "invariant") ;
ev->ListNext = (ParkEvent *) (w & ~LOCKBIT );
if (Atomic::cmpxchg_ptr (intptr_t(ev)|LOCKBIT, Lock, w) == w) break ;
}
while (ev->OnList != 0) {
ev->park() ;
}
}
}
// Release() must extract a successor from the list and then wake that thread.
// It can "pop" the front of the list or use a detach-modify-reattach (DMR) scheme
// similar to that used by ParkEvent::Allocate() and ::Release(). DMR-based
// Release() would :
// (A) CAS() or swap() null to *Lock, releasing the lock and detaching the list.
// (B) Extract a successor from the private list "in-hand"
// (C) attempt to CAS() the residual back into *Lock over null.
// If there were any newly arrived threads and the CAS() would fail.
// In that case Release() would detach the RATs, re-merge the list in-hand
// with the RATs and repeat as needed. Alternately, Release() might
// detach and extract a successor, but then pass the residual list to the wakee.
// The wakee would be responsible for reattaching and remerging before it
// competed for the lock.
//
// Both "pop" and DMR are immune from ABA corruption -- there can be
// multiple concurrent pushers, but only one popper or detacher.
// This implementation pops from the head of the list. This is unfair,
// but tends to provide excellent throughput as hot threads remain hot.
// (We wake recently run threads first).
void Thread::muxRelease (volatile intptr_t * Lock) {
for (;;) {
const intptr_t w = Atomic::cmpxchg_ptr (0, Lock, LOCKBIT) ;
assert (w & LOCKBIT, "invariant") ;
if (w == LOCKBIT) return ;
ParkEvent * List = (ParkEvent *) (w & ~LOCKBIT) ;
assert (List != NULL, "invariant") ;
assert (List->OnList == intptr_t(Lock), "invariant") ;
ParkEvent * nxt = List->ListNext ;
// The following CAS() releases the lock and pops the head element.
if (Atomic::cmpxchg_ptr (intptr_t(nxt), Lock, w) != w) {
continue ;
}
List->OnList = 0 ;
OrderAccess::fence() ;
List->unpark () ;
return ;
}
}
// ObjectMonitor Lifecycle
// -----------------------
// Inflation unlinks monitors from the global gFreeList and
// associates them with objects. Deflation -- which occurs at
// STW-time -- disassociates idle monitors from objects. Such
// scavenged monitors are returned to the gFreeList.
//
// The global list is protected by ListLock. All the critical sections
// are short and operate in constant-time.
//
// ObjectMonitors reside in type-stable memory (TSM) and are immortal.
//
// Lifecycle:
// -- unassigned and on the global free list
// -- unassigned and on a thread's private omFreeList
// -- assigned to an object. The object is inflated and the mark refers
// to the objectmonitor.
//
// TODO-FIXME:
//
// * We currently protect the gFreeList with a simple lock.
// An alternate lock-free scheme would be to pop elements from the gFreeList
// with CAS. This would be safe from ABA corruption as long we only
// recycled previously appearing elements onto the list in deflate_idle_monitors()
// at STW-time. Completely new elements could always be pushed onto the gFreeList
// with CAS. Elements that appeared previously on the list could only
// be installed at STW-time.
//
// * For efficiency and to help reduce the store-before-CAS penalty
// the objectmonitors on gFreeList or local free lists should be ready to install
// with the exception of _header and _object. _object can be set after inflation.
// In particular, keep all objectMonitors on a thread's private list in ready-to-install
// state with m.Owner set properly.
//
// * We could all diffuse contention by using multiple global (FreeList, Lock)
// pairs -- threads could use trylock() and a cyclic-scan strategy to search for
// an unlocked free list.
//
// * Add lifecycle tags and assert()s.
//
// * Be more consistent about when we clear an objectmonitor's fields:
// A. After extracting the objectmonitor from a free list.
// B. After adding an objectmonitor to a free list.
//
ObjectMonitor * ObjectSynchronizer::gBlockList = NULL ;
ObjectMonitor * volatile ObjectSynchronizer::gFreeList = NULL ;
static volatile intptr_t ListLock = 0 ; // protects global monitor free-list cache
#define CHAINMARKER ((oop)-1)
ObjectMonitor * ATTR ObjectSynchronizer::omAlloc (Thread * Self) {
// A large MAXPRIVATE value reduces both list lock contention
// and list coherency traffic, but also tends to increase the
// number of objectMonitors in circulation as well as the STW
// scavenge costs. As usual, we lean toward time in space-time
// tradeoffs.
const int MAXPRIVATE = 1024 ;
for (;;) {
ObjectMonitor * m ;
// 1: try to allocate from the thread's local omFreeList.
// Threads will attempt to allocate first from their local list, then
// from the global list, and only after those attempts fail will the thread
// attempt to instantiate new monitors. Thread-local free lists take
// heat off the ListLock and improve allocation latency, as well as reducing
// coherency traffic on the shared global list.
m = Self->omFreeList ;
if (m != NULL) {
Self->omFreeList = m->FreeNext ;
Self->omFreeCount -- ;
// CONSIDER: set m->FreeNext = BAD -- diagnostic hygiene
guarantee (m->object() == NULL, "invariant") ;
return m ;
}
// 2: try to allocate from the global gFreeList
// CONSIDER: use muxTry() instead of muxAcquire().
// If the muxTry() fails then drop immediately into case 3.
// If we're using thread-local free lists then try
// to reprovision the caller's free list.
if (gFreeList != NULL) {
// Reprovision the thread's omFreeList.
// Use bulk transfers to reduce the allocation rate and heat
// on various locks.
Thread::muxAcquire (&ListLock, "omAlloc") ;
for (int i = Self->omFreeProvision; --i >= 0 && gFreeList != NULL; ) {
ObjectMonitor * take = gFreeList ;
gFreeList = take->FreeNext ;
guarantee (take->object() == NULL, "invariant") ;
guarantee (!take->is_busy(), "invariant") ;
take->Recycle() ;
omRelease (Self, take) ;
}
Thread::muxRelease (&ListLock) ;
Self->omFreeProvision += 1 + (Self->omFreeProvision/2) ;
if (Self->omFreeProvision > MAXPRIVATE ) Self->omFreeProvision = MAXPRIVATE ;
TEVENT (omFirst - reprovision) ;
continue ;
}
// 3: allocate a block of new ObjectMonitors
// Both the local and global free lists are empty -- resort to malloc().
// In the current implementation objectMonitors are TSM - immortal.
assert (_BLOCKSIZE > 1, "invariant") ;
ObjectMonitor * temp = new ObjectMonitor[_BLOCKSIZE];
// NOTE: (almost) no way to recover if allocation failed.
// We might be able to induce a STW safepoint and scavenge enough
// objectMonitors to permit progress.
if (temp == NULL) {
vm_exit_out_of_memory (sizeof (ObjectMonitor[_BLOCKSIZE]), "Allocate ObjectMonitors") ;
}
// Format the block.
// initialize the linked list, each monitor points to its next
// forming the single linked free list, the very first monitor
// will points to next block, which forms the block list.
// The trick of using the 1st element in the block as gBlockList
// linkage should be reconsidered. A better implementation would
// look like: class Block { Block * next; int N; ObjectMonitor Body [N] ; }
for (int i = 1; i < _BLOCKSIZE ; i++) {
temp[i].FreeNext = &temp[i+1];
}
// terminate the last monitor as the end of list
temp[_BLOCKSIZE - 1].FreeNext = NULL ;
// Element [0] is reserved for global list linkage
temp[0].set_object(CHAINMARKER);
// Consider carving out this thread's current request from the
// block in hand. This avoids some lock traffic and redundant
// list activity.
// Acquire the ListLock to manipulate BlockList and FreeList.
// An Oyama-Taura-Yonezawa scheme might be more efficient.
Thread::muxAcquire (&ListLock, "omAlloc [2]") ;
// Add the new block to the list of extant blocks (gBlockList).
// The very first objectMonitor in a block is reserved and dedicated.
// It serves as blocklist "next" linkage.
temp[0].FreeNext = gBlockList;
gBlockList = temp;
// Add the new string of objectMonitors to the global free list
temp[_BLOCKSIZE - 1].FreeNext = gFreeList ;
gFreeList = temp + 1;
Thread::muxRelease (&ListLock) ;
TEVENT (Allocate block of monitors) ;
}
}
// Place "m" on the caller's private per-thread omFreeList.
// In practice there's no need to clamp or limit the number of
// monitors on a thread's omFreeList as the only time we'll call
// omRelease is to return a monitor to the free list after a CAS
// attempt failed. This doesn't allow unbounded #s of monitors to
// accumulate on a thread's free list.
//
// In the future the usage of omRelease() might change and monitors
// could migrate between free lists. In that case to avoid excessive
// accumulation we could limit omCount to (omProvision*2), otherwise return
// the objectMonitor to the global list. We should drain (return) in reasonable chunks.
// That is, *not* one-at-a-time.
void ObjectSynchronizer::omRelease (Thread * Self, ObjectMonitor * m) {
guarantee (m->object() == NULL, "invariant") ;
m->FreeNext = Self->omFreeList ;
Self->omFreeList = m ;
Self->omFreeCount ++ ;
}
// Return the monitors of a moribund thread's local free list to
// the global free list. Typically a thread calls omFlush() when
// it's dying. We could also consider having the VM thread steal
// monitors from threads that have not run java code over a few
// consecutive STW safepoints. Relatedly, we might decay
// omFreeProvision at STW safepoints.
//
// We currently call omFlush() from the Thread:: dtor _after the thread
// has been excised from the thread list and is no longer a mutator.
// That means that omFlush() can run concurrently with a safepoint and
// the scavenge operator. Calling omFlush() from JavaThread::exit() might
// be a better choice as we could safely reason that that the JVM is
// not at a safepoint at the time of the call, and thus there could
// be not inopportune interleavings between omFlush() and the scavenge
// operator.
void ObjectSynchronizer::omFlush (Thread * Self) {
ObjectMonitor * List = Self->omFreeList ; // Null-terminated SLL
Self->omFreeList = NULL ;
if (List == NULL) return ;
ObjectMonitor * Tail = NULL ;
ObjectMonitor * s ;
for (s = List ; s != NULL ; s = s->FreeNext) {
Tail = s ;
guarantee (s->object() == NULL, "invariant") ;
guarantee (!s->is_busy(), "invariant") ;
s->set_owner (NULL) ; // redundant but good hygiene
TEVENT (omFlush - Move one) ;
}
guarantee (Tail != NULL && List != NULL, "invariant") ;
Thread::muxAcquire (&ListLock, "omFlush") ;
Tail->FreeNext = gFreeList ;
gFreeList = List ;
Thread::muxRelease (&ListLock) ;
TEVENT (omFlush) ;
}
// Get the next block in the block list.
static inline ObjectMonitor* next(ObjectMonitor* block) {
assert(block->object() == CHAINMARKER, "must be a block header");
block = block->FreeNext ;
assert(block == NULL || block->object() == CHAINMARKER, "must be a block header");
return block;
}
// Fast path code shared by multiple functions
ObjectMonitor* ObjectSynchronizer::inflate_helper(oop obj) {
markOop mark = obj->mark();
if (mark->has_monitor()) {
assert(ObjectSynchronizer::verify_objmon_isinpool(mark->monitor()), "monitor is invalid");
assert(mark->monitor()->header()->is_neutral(), "monitor must record a good object header");
return mark->monitor();
}
return ObjectSynchronizer::inflate(Thread::current(), obj);
}
// Note that we could encounter some performance loss through false-sharing as
// multiple locks occupy the same $ line. Padding might be appropriate.
#define NINFLATIONLOCKS 256
static volatile intptr_t InflationLocks [NINFLATIONLOCKS] ;
static markOop ReadStableMark (oop obj) {
markOop mark = obj->mark() ;
if (!mark->is_being_inflated()) {
return mark ; // normal fast-path return
}
int its = 0 ;
for (;;) {
markOop mark = obj->mark() ;
if (!mark->is_being_inflated()) {
return mark ; // normal fast-path return
}
// The object is being inflated by some other thread.
// The caller of ReadStableMark() must wait for inflation to complete.
// Avoid live-lock
// TODO: consider calling SafepointSynchronize::do_call_back() while
// spinning to see if there's a safepoint pending. If so, immediately
// yielding or blocking would be appropriate. Avoid spinning while
// there is a safepoint pending.
// TODO: add inflation contention performance counters.
// TODO: restrict the aggregate number of spinners.
++its ;
if (its > 10000 || !os::is_MP()) {
if (its & 1) {
os::NakedYield() ;
TEVENT (Inflate: INFLATING - yield) ;
} else {
// Note that the following code attenuates the livelock problem but is not
// a complete remedy. A more complete solution would require that the inflating
// thread hold the associated inflation lock. The following code simply restricts
// the number of spinners to at most one. We'll have N-2 threads blocked
// on the inflationlock, 1 thread holding the inflation lock and using
// a yield/park strategy, and 1 thread in the midst of inflation.
// A more refined approach would be to change the encoding of INFLATING
// to allow encapsulation of a native thread pointer. Threads waiting for
// inflation to complete would use CAS to push themselves onto a singly linked
// list rooted at the markword. Once enqueued, they'd loop, checking a per-thread flag
// and calling park(). When inflation was complete the thread that accomplished inflation
// would detach the list and set the markword to inflated with a single CAS and
// then for each thread on the list, set the flag and unpark() the thread.
// This is conceptually similar to muxAcquire-muxRelease, except that muxRelease
// wakes at most one thread whereas we need to wake the entire list.
int ix = (intptr_t(obj) >> 5) & (NINFLATIONLOCKS-1) ;
int YieldThenBlock = 0 ;
assert (ix >= 0 && ix < NINFLATIONLOCKS, "invariant") ;
assert ((NINFLATIONLOCKS & (NINFLATIONLOCKS-1)) == 0, "invariant") ;
Thread::muxAcquire (InflationLocks + ix, "InflationLock") ;
while (obj->mark() == markOopDesc::INFLATING()) {
// Beware: NakedYield() is advisory and has almost no effect on some platforms
// so we periodically call Self->_ParkEvent->park(1).
// We use a mixed spin/yield/block mechanism.
if ((YieldThenBlock++) >= 16) {
Thread::current()->_ParkEvent->park(1) ;
} else {
os::NakedYield() ;
}
}
Thread::muxRelease (InflationLocks + ix ) ;
TEVENT (Inflate: INFLATING - yield/park) ;
}
} else {
SpinPause() ; // SMP-polite spinning
}
}
}
ObjectMonitor * ATTR ObjectSynchronizer::inflate (Thread * Self, oop object) {
// Inflate mutates the heap ...
// Relaxing assertion for bug 6320749.
assert (Universe::verify_in_progress() ||
!SafepointSynchronize::is_at_safepoint(), "invariant") ;
for (;;) {
const markOop mark = object->mark() ;
assert (!mark->has_bias_pattern(), "invariant") ;
// The mark can be in one of the following states:
// * Inflated - just return
// * Stack-locked - coerce it to inflated
// * INFLATING - busy wait for conversion to complete
// * Neutral - aggressively inflate the object.
// * BIASED - Illegal. We should never see this
// CASE: inflated
if (mark->has_monitor()) {
ObjectMonitor * inf = mark->monitor() ;
assert (inf->header()->is_neutral(), "invariant");
assert (inf->object() == object, "invariant") ;
assert (ObjectSynchronizer::verify_objmon_isinpool(inf), "monitor is invalid");
return inf ;
}
// CASE: inflation in progress - inflating over a stack-lock.
// Some other thread is converting from stack-locked to inflated.
// Only that thread can complete inflation -- other threads must wait.
// The INFLATING value is transient.
// Currently, we spin/yield/park and poll the markword, waiting for inflation to finish.
// We could always eliminate polling by parking the thread on some auxiliary list.
if (mark == markOopDesc::INFLATING()) {
TEVENT (Inflate: spin while INFLATING) ;
ReadStableMark(object) ;
continue ;
}
// CASE: stack-locked
// Could be stack-locked either by this thread or by some other thread.
//
// Note that we allocate the objectmonitor speculatively, _before_ attempting
// to install INFLATING into the mark word. We originally installed INFLATING,
// allocated the objectmonitor, and then finally STed the address of the
// objectmonitor into the mark. This was correct, but artificially lengthened
// the interval in which INFLATED appeared in the mark, thus increasing
// the odds of inflation contention.
//
// We now use per-thread private objectmonitor free lists.
// These list are reprovisioned from the global free list outside the
// critical INFLATING...ST interval. A thread can transfer
// multiple objectmonitors en-mass from the global free list to its local free list.
// This reduces coherency traffic and lock contention on the global free list.
// Using such local free lists, it doesn't matter if the omAlloc() call appears
// before or after the CAS(INFLATING) operation.
// See the comments in omAlloc().
if (mark->has_locker()) {
ObjectMonitor * m = omAlloc (Self) ;
// Optimistically prepare the objectmonitor - anticipate successful CAS
// We do this before the CAS in order to minimize the length of time
// in which INFLATING appears in the mark.
m->Recycle();
m->FreeNext = NULL ;
m->_Responsible = NULL ;
m->OwnerIsThread = 0 ;
m->_recursions = 0 ;
m->_SpinDuration = Knob_SpinLimit ; // Consider: maintain by type/class
markOop cmp = (markOop) Atomic::cmpxchg_ptr (markOopDesc::INFLATING(), object->mark_addr(), mark) ;
if (cmp != mark) {
omRelease (Self, m) ;
continue ; // Interference -- just retry
}
// We've successfully installed INFLATING (0) into the mark-word.
// This is the only case where 0 will appear in a mark-work.
// Only the singular thread that successfully swings the mark-word
// to 0 can perform (or more precisely, complete) inflation.
//
// Why do we CAS a 0 into the mark-word instead of just CASing the
// mark-word from the stack-locked value directly to the new inflated state?
// Consider what happens when a thread unlocks a stack-locked object.
// It attempts to use CAS to swing the displaced header value from the
// on-stack basiclock back into the object header. Recall also that the
// header value (hashcode, etc) can reside in (a) the object header, or
// (b) a displaced header associated with the stack-lock, or (c) a displaced
// header in an objectMonitor. The inflate() routine must copy the header
// value from the basiclock on the owner's stack to the objectMonitor, all
// the while preserving the hashCode stability invariants. If the owner
// decides to release the lock while the value is 0, the unlock will fail
// and control will eventually pass from slow_exit() to inflate. The owner
// will then spin, waiting for the 0 value to disappear. Put another way,
// the 0 causes the owner to stall if the owner happens to try to
// drop the lock (restoring the header from the basiclock to the object)
// while inflation is in-progress. This protocol avoids races that might
// would otherwise permit hashCode values to change or "flicker" for an object.
// Critically, while object->mark is 0 mark->displaced_mark_helper() is stable.
// 0 serves as a "BUSY" inflate-in-progress indicator.
// fetch the displaced mark from the owner's stack.
// The owner can't die or unwind past the lock while our INFLATING
// object is in the mark. Furthermore the owner can't complete
// an unlock on the object, either.
markOop dmw = mark->displaced_mark_helper() ;
assert (dmw->is_neutral(), "invariant") ;
// Setup monitor fields to proper values -- prepare the monitor
m->set_header(dmw) ;
// Optimization: if the mark->locker stack address is associated
// with this thread we could simply set m->_owner = Self and
// m->OwnerIsThread = 1. Note that a thread can inflate an object
// that it has stack-locked -- as might happen in wait() -- directly
// with CAS. That is, we can avoid the xchg-NULL .... ST idiom.
m->set_owner(mark->locker());
m->set_object(object);
// TODO-FIXME: assert BasicLock->dhw != 0.
// Must preserve store ordering. The monitor state must
// be stable at the time of publishing the monitor address.
guarantee (object->mark() == markOopDesc::INFLATING(), "invariant") ;
object->release_set_mark(markOopDesc::encode(m));
// Hopefully the performance counters are allocated on distinct cache lines
// to avoid false sharing on MP systems ...
if (_sync_Inflations != NULL) _sync_Inflations->inc() ;
TEVENT(Inflate: overwrite stacklock) ;
if (TraceMonitorInflation) {
if (object->is_instance()) {
ResourceMark rm;
tty->print_cr("Inflating object " INTPTR_FORMAT " , mark " INTPTR_FORMAT " , type %s",
(intptr_t) object, (intptr_t) object->mark(),
Klass::cast(object->klass())->external_name());
}
}
return m ;
}
// CASE: neutral
// TODO-FIXME: for entry we currently inflate and then try to CAS _owner.
// If we know we're inflating for entry it's better to inflate by swinging a
// pre-locked objectMonitor pointer into the object header. A successful
// CAS inflates the object *and* confers ownership to the inflating thread.
// In the current implementation we use a 2-step mechanism where we CAS()
// to inflate and then CAS() again to try to swing _owner from NULL to Self.
// An inflateTry() method that we could call from fast_enter() and slow_enter()
// would be useful.
assert (mark->is_neutral(), "invariant");
ObjectMonitor * m = omAlloc (Self) ;
// prepare m for installation - set monitor to initial state
m->Recycle();
m->set_header(mark);
m->set_owner(NULL);
m->set_object(object);
m->OwnerIsThread = 1 ;
m->_recursions = 0 ;
m->FreeNext = NULL ;
m->_Responsible = NULL ;
m->_SpinDuration = Knob_SpinLimit ; // consider: keep metastats by type/class
if (Atomic::cmpxchg_ptr (markOopDesc::encode(m), object->mark_addr(), mark) != mark) {
m->set_object (NULL) ;
m->set_owner (NULL) ;
m->OwnerIsThread = 0 ;
m->Recycle() ;
omRelease (Self, m) ;
m = NULL ;
continue ;
// interference - the markword changed - just retry.
// The state-transitions are one-way, so there's no chance of
// live-lock -- "Inflated" is an absorbing state.
}
// Hopefully the performance counters are allocated on distinct
// cache lines to avoid false sharing on MP systems ...
if (_sync_Inflations != NULL) _sync_Inflations->inc() ;
TEVENT(Inflate: overwrite neutral) ;
if (TraceMonitorInflation) {
if (object->is_instance()) {
ResourceMark rm;
tty->print_cr("Inflating object " INTPTR_FORMAT " , mark " INTPTR_FORMAT " , type %s",
(intptr_t) object, (intptr_t) object->mark(),
Klass::cast(object->klass())->external_name());
}
}
return m ;
}
}
// This the fast monitor enter. The interpreter and compiler use
// some assembly copies of this code. Make sure update those code
// if the following function is changed. The implementation is
// extremely sensitive to race condition. Be careful.
void ObjectSynchronizer::fast_enter(Handle obj, BasicLock* lock, bool attempt_rebias, TRAPS) {
if (UseBiasedLocking) {
if (!SafepointSynchronize::is_at_safepoint()) {
BiasedLocking::Condition cond = BiasedLocking::revoke_and_rebias(obj, attempt_rebias, THREAD);
if (cond == BiasedLocking::BIAS_REVOKED_AND_REBIASED) {
return;
}
} else {
assert(!attempt_rebias, "can not rebias toward VM thread");
BiasedLocking::revoke_at_safepoint(obj);
}
assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
}
slow_enter (obj, lock, THREAD) ;
}
void ObjectSynchronizer::fast_exit(oop object, BasicLock* lock, TRAPS) {
assert(!object->mark()->has_bias_pattern(), "should not see bias pattern here");
// if displaced header is null, the previous enter is recursive enter, no-op
markOop dhw = lock->displaced_header();
markOop mark ;
if (dhw == NULL) {
// Recursive stack-lock.
// Diagnostics -- Could be: stack-locked, inflating, inflated.
mark = object->mark() ;
assert (!mark->is_neutral(), "invariant") ;
if (mark->has_locker() && mark != markOopDesc::INFLATING()) {
assert(THREAD->is_lock_owned((address)mark->locker()), "invariant") ;
}
if (mark->has_monitor()) {
ObjectMonitor * m = mark->monitor() ;
assert(((oop)(m->object()))->mark() == mark, "invariant") ;
assert(m->is_entered(THREAD), "invariant") ;
}
return ;
}
mark = object->mark() ;
// If the object is stack-locked by the current thread, try to
// swing the displaced header from the box back to the mark.
if (mark == (markOop) lock) {
assert (dhw->is_neutral(), "invariant") ;
if ((markOop) Atomic::cmpxchg_ptr (dhw, object->mark_addr(), mark) == mark) {
TEVENT (fast_exit: release stacklock) ;
return;
}
}
ObjectSynchronizer::inflate(THREAD, object)->exit (THREAD) ;
}
// This routine is used to handle interpreter/compiler slow case
// We don't need to use fast path here, because it must have been
// failed in the interpreter/compiler code.
void ObjectSynchronizer::slow_enter(Handle obj, BasicLock* lock, TRAPS) {
markOop mark = obj->mark();
assert(!mark->has_bias_pattern(), "should not see bias pattern here");
if (mark->is_neutral()) {
// Anticipate successful CAS -- the ST of the displaced mark must
// be visible <= the ST performed by the CAS.
lock->set_displaced_header(mark);
if (mark == (markOop) Atomic::cmpxchg_ptr(lock, obj()->mark_addr(), mark)) {
TEVENT (slow_enter: release stacklock) ;
return ;
}
// Fall through to inflate() ...
} else
if (mark->has_locker() && THREAD->is_lock_owned((address)mark->locker())) {
assert(lock != mark->locker(), "must not re-lock the same lock");
assert(lock != (BasicLock*)obj->mark(), "don't relock with same BasicLock");
lock->set_displaced_header(NULL);
return;
}
#if 0
// The following optimization isn't particularly useful.
if (mark->has_monitor() && mark->monitor()->is_entered(THREAD)) {
lock->set_displaced_header (NULL) ;
return ;
}
#endif
// The object header will never be displaced to this lock,
// so it does not matter what the value is, except that it
// must be non-zero to avoid looking like a re-entrant lock,
// and must not look locked either.
lock->set_displaced_header(markOopDesc::unused_mark());
ObjectSynchronizer::inflate(THREAD, obj())->enter(THREAD);
}
// This routine is used to handle interpreter/compiler slow case
// We don't need to use fast path here, because it must have
// failed in the interpreter/compiler code. Simply use the heavy
// weight monitor should be ok, unless someone find otherwise.
void ObjectSynchronizer::slow_exit(oop object, BasicLock* lock, TRAPS) {
fast_exit (object, lock, THREAD) ;
}
// NOTE: must use heavy weight monitor to handle jni monitor enter
void ObjectSynchronizer::jni_enter(Handle obj, TRAPS) { // possible entry from jni enter
// the current locking is from JNI instead of Java code
TEVENT (jni_enter) ;
if (UseBiasedLocking) {
BiasedLocking::revoke_and_rebias(obj, false, THREAD);
assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
}
THREAD->set_current_pending_monitor_is_from_java(false);
ObjectSynchronizer::inflate(THREAD, obj())->enter(THREAD);
THREAD->set_current_pending_monitor_is_from_java(true);
}
// NOTE: must use heavy weight monitor to handle jni monitor enter
bool ObjectSynchronizer::jni_try_enter(Handle obj, Thread* THREAD) {
if (UseBiasedLocking) {
BiasedLocking::revoke_and_rebias(obj, false, THREAD);
assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
}
ObjectMonitor* monitor = ObjectSynchronizer::inflate_helper(obj());
return monitor->try_enter(THREAD);
}
// NOTE: must use heavy weight monitor to handle jni monitor exit
void ObjectSynchronizer::jni_exit(oop obj, Thread* THREAD) {
TEVENT (jni_exit) ;
if (UseBiasedLocking) {
BiasedLocking::revoke_and_rebias(obj, false, THREAD);
}
assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
ObjectMonitor* monitor = ObjectSynchronizer::inflate(THREAD, obj);
// If this thread has locked the object, exit the monitor. Note: can't use
// monitor->check(CHECK); must exit even if an exception is pending.
if (monitor->check(THREAD)) {
monitor->exit(THREAD);
}
}
// complete_exit()/reenter() are used to wait on a nested lock
// i.e. to give up an outer lock completely and then re-enter
// Used when holding nested locks - lock acquisition order: lock1 then lock2
// 1) complete_exit lock1 - saving recursion count
// 2) wait on lock2
// 3) when notified on lock2, unlock lock2
// 4) reenter lock1 with original recursion count
// 5) lock lock2
// NOTE: must use heavy weight monitor to handle complete_exit/reenter()
intptr_t ObjectSynchronizer::complete_exit(Handle obj, TRAPS) {
TEVENT (complete_exit) ;
if (UseBiasedLocking) {
BiasedLocking::revoke_and_rebias(obj, false, THREAD);
assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
}
ObjectMonitor* monitor = ObjectSynchronizer::inflate(THREAD, obj());
return monitor->complete_exit(THREAD);
}
// NOTE: must use heavy weight monitor to handle complete_exit/reenter()
void ObjectSynchronizer::reenter(Handle obj, intptr_t recursion, TRAPS) {
TEVENT (reenter) ;
if (UseBiasedLocking) {
BiasedLocking::revoke_and_rebias(obj, false, THREAD);
assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
}
ObjectMonitor* monitor = ObjectSynchronizer::inflate(THREAD, obj());
monitor->reenter(recursion, THREAD);
}
// This exists only as a workaround of dtrace bug 6254741
int dtrace_waited_probe(ObjectMonitor* monitor, Handle obj, Thread* thr) {
DTRACE_MONITOR_PROBE(waited, monitor, obj(), thr);
return 0;
}
// NOTE: must use heavy weight monitor to handle wait()
void ObjectSynchronizer::wait(Handle obj, jlong millis, TRAPS) {
if (UseBiasedLocking) {
BiasedLocking::revoke_and_rebias(obj, false, THREAD);
assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
}
if (millis < 0) {
TEVENT (wait - throw IAX) ;
THROW_MSG(vmSymbols::java_lang_IllegalArgumentException(), "timeout value is negative");
}
ObjectMonitor* monitor = ObjectSynchronizer::inflate(THREAD, obj());
DTRACE_MONITOR_WAIT_PROBE(monitor, obj(), THREAD, millis);
monitor->wait(millis, true, THREAD);
/* This dummy call is in place to get around dtrace bug 6254741. Once
that's fixed we can uncomment the following line and remove the call */
// DTRACE_MONITOR_PROBE(waited, monitor, obj(), THREAD);
dtrace_waited_probe(monitor, obj, THREAD);
}
void ObjectSynchronizer::waitUninterruptibly (Handle obj, jlong millis, TRAPS) {
if (UseBiasedLocking) {
BiasedLocking::revoke_and_rebias(obj, false, THREAD);
assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
}
if (millis < 0) {
TEVENT (wait - throw IAX) ;
THROW_MSG(vmSymbols::java_lang_IllegalArgumentException(), "timeout value is negative");
}
ObjectSynchronizer::inflate(THREAD, obj()) -> wait(millis, false, THREAD) ;
}
void ObjectSynchronizer::notify(Handle obj, TRAPS) {
if (UseBiasedLocking) {
BiasedLocking::revoke_and_rebias(obj, false, THREAD);
assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
}
markOop mark = obj->mark();
if (mark->has_locker() && THREAD->is_lock_owned((address)mark->locker())) {
return;
}
ObjectSynchronizer::inflate(THREAD, obj())->notify(THREAD);
}
// NOTE: see comment of notify()
void ObjectSynchronizer::notifyall(Handle obj, TRAPS) {
if (UseBiasedLocking) {
BiasedLocking::revoke_and_rebias(obj, false, THREAD);
assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
}
markOop mark = obj->mark();
if (mark->has_locker() && THREAD->is_lock_owned((address)mark->locker())) {
return;
}
ObjectSynchronizer::inflate(THREAD, obj())->notifyAll(THREAD);
}
intptr_t ObjectSynchronizer::FastHashCode (Thread * Self, oop obj) {
if (UseBiasedLocking) {
// NOTE: many places throughout the JVM do not expect a safepoint
// to be taken here, in particular most operations on perm gen
// objects. However, we only ever bias Java instances and all of
// the call sites of identity_hash that might revoke biases have
// been checked to make sure they can handle a safepoint. The
// added check of the bias pattern is to avoid useless calls to
// thread-local storage.
if (obj->mark()->has_bias_pattern()) {
// Box and unbox the raw reference just in case we cause a STW safepoint.
Handle hobj (Self, obj) ;
// Relaxing assertion for bug 6320749.
assert (Universe::verify_in_progress() ||
!SafepointSynchronize::is_at_safepoint(),
"biases should not be seen by VM thread here");
BiasedLocking::revoke_and_rebias(hobj, false, JavaThread::current());
obj = hobj() ;
assert(!obj->mark()->has_bias_pattern(), "biases should be revoked by now");
}
}
// hashCode() is a heap mutator ...
// Relaxing assertion for bug 6320749.
assert (Universe::verify_in_progress() ||
!SafepointSynchronize::is_at_safepoint(), "invariant") ;
assert (Universe::verify_in_progress() ||
Self->is_Java_thread() , "invariant") ;
assert (Universe::verify_in_progress() ||
((JavaThread *)Self)->thread_state() != _thread_blocked, "invariant") ;
ObjectMonitor* monitor = NULL;
markOop temp, test;
intptr_t hash;
markOop mark = ReadStableMark (obj);
// object should remain ineligible for biased locking
assert (!mark->has_bias_pattern(), "invariant") ;
if (mark->is_neutral()) {
hash = mark->hash(); // this is a normal header
if (hash) { // if it has hash, just return it
return hash;
}
hash = get_next_hash(Self, obj); // allocate a new hash code
temp = mark->copy_set_hash(hash); // merge the hash code into header
// use (machine word version) atomic operation to install the hash
test = (markOop) Atomic::cmpxchg_ptr(temp, obj->mark_addr(), mark);
if (test == mark) {
return hash;
}
// If atomic operation failed, we must inflate the header
// into heavy weight monitor. We could add more code here
// for fast path, but it does not worth the complexity.
} else if (mark->has_monitor()) {
monitor = mark->monitor();
temp = monitor->header();
assert (temp->is_neutral(), "invariant") ;
hash = temp->hash();
if (hash) {
return hash;
}
// Skip to the following code to reduce code size
} else if (Self->is_lock_owned((address)mark->locker())) {
temp = mark->displaced_mark_helper(); // this is a lightweight monitor owned
assert (temp->is_neutral(), "invariant") ;
hash = temp->hash(); // by current thread, check if the displaced
if (hash) { // header contains hash code
return hash;
}
// WARNING:
// The displaced header is strictly immutable.
// It can NOT be changed in ANY cases. So we have
// to inflate the header into heavyweight monitor
// even the current thread owns the lock. The reason
// is the BasicLock (stack slot) will be asynchronously
// read by other threads during the inflate() function.
// Any change to stack may not propagate to other threads
// correctly.
}
// Inflate the monitor to set hash code
monitor = ObjectSynchronizer::inflate(Self, obj);
// Load displaced header and check it has hash code
mark = monitor->header();
assert (mark->is_neutral(), "invariant") ;
hash = mark->hash();
if (hash == 0) {
hash = get_next_hash(Self, obj);
temp = mark->copy_set_hash(hash); // merge hash code into header
assert (temp->is_neutral(), "invariant") ;
test = (markOop) Atomic::cmpxchg_ptr(temp, monitor, mark);
if (test != mark) {
// The only update to the header in the monitor (outside GC)
// is install the hash code. If someone add new usage of
// displaced header, please update this code
hash = test->hash();
assert (test->is_neutral(), "invariant") ;
assert (hash != 0, "Trivial unexpected object/monitor header usage.");
}
}
// We finally get the hash
return hash;
}
// Deprecated -- use FastHashCode() instead.
intptr_t ObjectSynchronizer::identity_hash_value_for(Handle obj) {
return FastHashCode (Thread::current(), obj()) ;
}
bool ObjectSynchronizer::current_thread_holds_lock(JavaThread* thread,
Handle h_obj) {
if (UseBiasedLocking) {
BiasedLocking::revoke_and_rebias(h_obj, false, thread);
assert(!h_obj->mark()->has_bias_pattern(), "biases should be revoked by now");
}
assert(thread == JavaThread::current(), "Can only be called on current thread");
oop obj = h_obj();
markOop mark = ReadStableMark (obj) ;
// Uncontended case, header points to stack
if (mark->has_locker()) {
return thread->is_lock_owned((address)mark->locker());
}
// Contended case, header points to ObjectMonitor (tagged pointer)
if (mark->has_monitor()) {
ObjectMonitor* monitor = mark->monitor();
return monitor->is_entered(thread) != 0 ;
}
// Unlocked case, header in place
assert(mark->is_neutral(), "sanity check");
return false;
}
// Be aware of this method could revoke bias of the lock object.
// This method querys the ownership of the lock handle specified by 'h_obj'.
// If the current thread owns the lock, it returns owner_self. If no
// thread owns the lock, it returns owner_none. Otherwise, it will return
// ower_other.
ObjectSynchronizer::LockOwnership ObjectSynchronizer::query_lock_ownership
(JavaThread *self, Handle h_obj) {
// The caller must beware this method can revoke bias, and
// revocation can result in a safepoint.
assert (!SafepointSynchronize::is_at_safepoint(), "invariant") ;
assert (self->thread_state() != _thread_blocked , "invariant") ;
// Possible mark states: neutral, biased, stack-locked, inflated
if (UseBiasedLocking && h_obj()->mark()->has_bias_pattern()) {
// CASE: biased
BiasedLocking::revoke_and_rebias(h_obj, false, self);
assert(!h_obj->mark()->has_bias_pattern(),
"biases should be revoked by now");
}
assert(self == JavaThread::current(), "Can only be called on current thread");
oop obj = h_obj();
markOop mark = ReadStableMark (obj) ;
// CASE: stack-locked. Mark points to a BasicLock on the owner's stack.
if (mark->has_locker()) {
return self->is_lock_owned((address)mark->locker()) ?
owner_self : owner_other;
}
// CASE: inflated. Mark (tagged pointer) points to an objectMonitor.
// The Object:ObjectMonitor relationship is stable as long as we're
// not at a safepoint.
if (mark->has_monitor()) {
void * owner = mark->monitor()->_owner ;
if (owner == NULL) return owner_none ;
return (owner == self ||
self->is_lock_owned((address)owner)) ? owner_self : owner_other;
}
// CASE: neutral
assert(mark->is_neutral(), "sanity check");
return owner_none ; // it's unlocked
}
// FIXME: jvmti should call this
JavaThread* ObjectSynchronizer::get_lock_owner(Handle h_obj, bool doLock) {
if (UseBiasedLocking) {
if (SafepointSynchronize::is_at_safepoint()) {
BiasedLocking::revoke_at_safepoint(h_obj);
} else {
BiasedLocking::revoke_and_rebias(h_obj, false, JavaThread::current());
}
assert(!h_obj->mark()->has_bias_pattern(), "biases should be revoked by now");
}
oop obj = h_obj();
address owner = NULL;
markOop mark = ReadStableMark (obj) ;
// Uncontended case, header points to stack
if (mark->has_locker()) {
owner = (address) mark->locker();
}
// Contended case, header points to ObjectMonitor (tagged pointer)
if (mark->has_monitor()) {
ObjectMonitor* monitor = mark->monitor();
assert(monitor != NULL, "monitor should be non-null");
owner = (address) monitor->owner();
}
if (owner != NULL) {
return Threads::owning_thread_from_monitor_owner(owner, doLock);
}
// Unlocked case, header in place
// Cannot have assertion since this object may have been
// locked by another thread when reaching here.
// assert(mark->is_neutral(), "sanity check");
return NULL;
}
// Iterate through monitor cache and attempt to release thread's monitors
// Gives up on a particular monitor if an exception occurs, but continues
// the overall iteration, swallowing the exception.
class ReleaseJavaMonitorsClosure: public MonitorClosure {
private:
TRAPS;
public:
ReleaseJavaMonitorsClosure(Thread* thread) : THREAD(thread) {}
void do_monitor(ObjectMonitor* mid) {
if (mid->owner() == THREAD) {
(void)mid->complete_exit(CHECK);
}
}
};
// Release all inflated monitors owned by THREAD. Lightweight monitors are
// ignored. This is meant to be called during JNI thread detach which assumes
// all remaining monitors are heavyweight. All exceptions are swallowed.
// Scanning the extant monitor list can be time consuming.
// A simple optimization is to add a per-thread flag that indicates a thread
// called jni_monitorenter() during its lifetime.
//
// Instead of No_Savepoint_Verifier it might be cheaper to
// use an idiom of the form:
// auto int tmp = SafepointSynchronize::_safepoint_counter ;
// <code that must not run at safepoint>
// guarantee (((tmp ^ _safepoint_counter) | (tmp & 1)) == 0) ;
// Since the tests are extremely cheap we could leave them enabled
// for normal product builds.
void ObjectSynchronizer::release_monitors_owned_by_thread(TRAPS) {
assert(THREAD == JavaThread::current(), "must be current Java thread");
No_Safepoint_Verifier nsv ;
ReleaseJavaMonitorsClosure rjmc(THREAD);
Thread::muxAcquire(&ListLock, "release_monitors_owned_by_thread");
ObjectSynchronizer::monitors_iterate(&rjmc);
Thread::muxRelease(&ListLock);
THREAD->clear_pending_exception();
}
// Visitors ...
void ObjectSynchronizer::monitors_iterate(MonitorClosure* closure) {
ObjectMonitor* block = gBlockList;
ObjectMonitor* mid;
while (block) {
assert(block->object() == CHAINMARKER, "must be a block header");
for (int i = _BLOCKSIZE - 1; i > 0; i--) {
mid = block + i;
oop object = (oop) mid->object();
if (object != NULL) {
closure->do_monitor(mid);
}
}
block = (ObjectMonitor*) block->FreeNext;
}
}
void ObjectSynchronizer::oops_do(OopClosure* f) {
assert(SafepointSynchronize::is_at_safepoint(), "must be at safepoint");
for (ObjectMonitor* block = gBlockList; block != NULL; block = next(block)) {
assert(block->object() == CHAINMARKER, "must be a block header");
for (int i = 1; i < _BLOCKSIZE; i++) {
ObjectMonitor* mid = &block[i];
if (mid->object() != NULL) {
f->do_oop((oop*)mid->object_addr());
}
}
}
}
// Deflate_idle_monitors() is called at all safepoints, immediately
// after all mutators are stopped, but before any objects have moved.
// It traverses the list of known monitors, deflating where possible.
// The scavenged monitor are returned to the monitor free list.
//
// Beware that we scavenge at *every* stop-the-world point.
// Having a large number of monitors in-circulation negatively
// impacts the performance of some applications (e.g., PointBase).
// Broadly, we want to minimize the # of monitors in circulation.
// Alternately, we could partition the active monitors into sub-lists
// of those that need scanning and those that do not.
// Specifically, we would add a new sub-list of objectmonitors
// that are in-circulation and potentially active. deflate_idle_monitors()
// would scan only that list. Other monitors could reside on a quiescent
// list. Such sequestered monitors wouldn't need to be scanned by
// deflate_idle_monitors(). omAlloc() would first check the global free list,
// then the quiescent list, and, failing those, would allocate a new block.
// Deflate_idle_monitors() would scavenge and move monitors to the
// quiescent list.
//
// Perversely, the heap size -- and thus the STW safepoint rate --
// typically drives the scavenge rate. Large heaps can mean infrequent GC,
// which in turn can mean large(r) numbers of objectmonitors in circulation.
// This is an unfortunate aspect of this design.
//
// Another refinement would be to refrain from calling deflate_idle_monitors()
// except at stop-the-world points associated with garbage collections.
//
// An even better solution would be to deflate on-the-fly, aggressively,
// at monitorexit-time as is done in EVM's metalock or Relaxed Locks.
void ObjectSynchronizer::deflate_idle_monitors() {
assert(SafepointSynchronize::is_at_safepoint(), "must be at safepoint");
int nInuse = 0 ; // currently associated with objects
int nInCirculation = 0 ; // extant
int nScavenged = 0 ; // reclaimed
ObjectMonitor * FreeHead = NULL ; // Local SLL of scavenged monitors
ObjectMonitor * FreeTail = NULL ;
// Iterate over all extant monitors - Scavenge all idle monitors.
TEVENT (deflate_idle_monitors) ;
for (ObjectMonitor* block = gBlockList; block != NULL; block = next(block)) {
assert(block->object() == CHAINMARKER, "must be a block header");
nInCirculation += _BLOCKSIZE ;
for (int i = 1 ; i < _BLOCKSIZE; i++) {
ObjectMonitor* mid = &block[i];
oop obj = (oop) mid->object();
if (obj == NULL) {
// The monitor is not associated with an object.
// The monitor should either be a thread-specific private
// free list or the global free list.
// obj == NULL IMPLIES mid->is_busy() == 0
guarantee (!mid->is_busy(), "invariant") ;
continue ;
}
// Normal case ... The monitor is associated with obj.
guarantee (obj->mark() == markOopDesc::encode(mid), "invariant") ;
guarantee (mid == obj->mark()->monitor(), "invariant");
guarantee (mid->header()->is_neutral(), "invariant");
if (mid->is_busy()) {
if (ClearResponsibleAtSTW) mid->_Responsible = NULL ;
nInuse ++ ;
} else {
// Deflate the monitor if it is no longer being used
// It's idle - scavenge and return to the global free list
// plain old deflation ...
TEVENT (deflate_idle_monitors - scavenge1) ;
if (TraceMonitorInflation) {
if (obj->is_instance()) {
ResourceMark rm;
tty->print_cr("Deflating object " INTPTR_FORMAT " , mark " INTPTR_FORMAT " , type %s",
(intptr_t) obj, (intptr_t) obj->mark(), Klass::cast(obj->klass())->external_name());
}
}
// Restore the header back to obj
obj->release_set_mark(mid->header());
mid->clear();
assert (mid->object() == NULL, "invariant") ;
// Move the object to the working free list defined by FreeHead,FreeTail.
mid->FreeNext = NULL ;
if (FreeHead == NULL) FreeHead = mid ;
if (FreeTail != NULL) FreeTail->FreeNext = mid ;
FreeTail = mid ;
nScavenged ++ ;
}
}
}
// Move the scavenged monitors back to the global free list.
// In theory we don't need the freelist lock as we're at a STW safepoint.
// omAlloc() and omFree() can only be called while a thread is _not in safepoint state.
// But it's remotely possible that omFlush() or release_monitors_owned_by_thread()
// might be called while not at a global STW safepoint. In the interest of
// safety we protect the following access with ListLock.
// An even more conservative and prudent approach would be to guard
// the main loop in scavenge_idle_monitors() with ListLock.
if (FreeHead != NULL) {
guarantee (FreeTail != NULL && nScavenged > 0, "invariant") ;
assert (FreeTail->FreeNext == NULL, "invariant") ;
// constant-time list splice - prepend scavenged segment to gFreeList
Thread::muxAcquire (&ListLock, "scavenge - return") ;
FreeTail->FreeNext = gFreeList ;
gFreeList = FreeHead ;
Thread::muxRelease (&ListLock) ;
}
if (_sync_Deflations != NULL) _sync_Deflations->inc(nScavenged) ;
if (_sync_MonExtant != NULL) _sync_MonExtant ->set_value(nInCirculation);
// TODO: Add objectMonitor leak detection.
// Audit/inventory the objectMonitors -- make sure they're all accounted for.
GVars.stwRandom = os::random() ;
GVars.stwCycle ++ ;
}
// A macro is used below because there may already be a pending
// exception which should not abort the execution of the routines
// which use this (which is why we don't put this into check_slow and
// call it with a CHECK argument).
#define CHECK_OWNER() \
do { \
if (THREAD != _owner) { \
if (THREAD->is_lock_owned((address) _owner)) { \
_owner = THREAD ; /* Convert from basiclock addr to Thread addr */ \
_recursions = 0; \
OwnerIsThread = 1 ; \
} else { \
TEVENT (Throw IMSX) ; \
THROW(vmSymbols::java_lang_IllegalMonitorStateException()); \
} \
} \
} while (false)
// TODO-FIXME: eliminate ObjectWaiters. Replace this visitor/enumerator
// interface with a simple FirstWaitingThread(), NextWaitingThread() interface.
ObjectWaiter* ObjectMonitor::first_waiter() {
return _WaitSet;
}
ObjectWaiter* ObjectMonitor::next_waiter(ObjectWaiter* o) {
return o->_next;
}
Thread* ObjectMonitor::thread_of_waiter(ObjectWaiter* o) {
return o->_thread;
}
// initialize the monitor, exception the semaphore, all other fields
// are simple integers or pointers
ObjectMonitor::ObjectMonitor() {
_header = NULL;
_count = 0;
_waiters = 0,
_recursions = 0;
_object = NULL;
_owner = NULL;
_WaitSet = NULL;
_WaitSetLock = 0 ;
_Responsible = NULL ;
_succ = NULL ;
_cxq = NULL ;
FreeNext = NULL ;
_EntryList = NULL ;
_SpinFreq = 0 ;
_SpinClock = 0 ;
OwnerIsThread = 0 ;
}
ObjectMonitor::~ObjectMonitor() {
// TODO: Add asserts ...
// _cxq == 0 _succ == NULL _owner == NULL _waiters == 0
// _count == 0 _EntryList == NULL etc
}
intptr_t ObjectMonitor::is_busy() const {
// TODO-FIXME: merge _count and _waiters.
// TODO-FIXME: assert _owner == null implies _recursions = 0
// TODO-FIXME: assert _WaitSet != null implies _count > 0
return _count|_waiters|intptr_t(_owner)|intptr_t(_cxq)|intptr_t(_EntryList ) ;
}
void ObjectMonitor::Recycle () {
// TODO: add stronger asserts ...
// _cxq == 0 _succ == NULL _owner == NULL _waiters == 0
// _count == 0 EntryList == NULL
// _recursions == 0 _WaitSet == NULL
// TODO: assert (is_busy()|_recursions) == 0
_succ = NULL ;
_EntryList = NULL ;
_cxq = NULL ;
_WaitSet = NULL ;
_recursions = 0 ;
_SpinFreq = 0 ;
_SpinClock = 0 ;
OwnerIsThread = 0 ;
}
// WaitSet management ...
inline void ObjectMonitor::AddWaiter(ObjectWaiter* node) {
assert(node != NULL, "should not dequeue NULL node");
assert(node->_prev == NULL, "node already in list");
assert(node->_next == NULL, "node already in list");
// put node at end of queue (circular doubly linked list)
if (_WaitSet == NULL) {
_WaitSet = node;
node->_prev = node;
node->_next = node;
} else {
ObjectWaiter* head = _WaitSet ;
ObjectWaiter* tail = head->_prev;
assert(tail->_next == head, "invariant check");
tail->_next = node;
head->_prev = node;
node->_next = head;
node->_prev = tail;
}
}
inline ObjectWaiter* ObjectMonitor::DequeueWaiter() {
// dequeue the very first waiter
ObjectWaiter* waiter = _WaitSet;
if (waiter) {
DequeueSpecificWaiter(waiter);
}
return waiter;
}
inline void ObjectMonitor::DequeueSpecificWaiter(ObjectWaiter* node) {
assert(node != NULL, "should not dequeue NULL node");
assert(node->_prev != NULL, "node already removed from list");
assert(node->_next != NULL, "node already removed from list");
// when the waiter has woken up because of interrupt,
// timeout or other spurious wake-up, dequeue the
// waiter from waiting list
ObjectWaiter* next = node->_next;
if (next == node) {
assert(node->_prev == node, "invariant check");
_WaitSet = NULL;
} else {
ObjectWaiter* prev = node->_prev;
assert(prev->_next == node, "invariant check");
assert(next->_prev == node, "invariant check");
next->_prev = prev;
prev->_next = next;
if (_WaitSet == node) {
_WaitSet = next;
}
}
node->_next = NULL;
node->_prev = NULL;
}
static char * kvGet (char * kvList, const char * Key) {
if (kvList == NULL) return NULL ;
size_t n = strlen (Key) ;
char * Search ;
for (Search = kvList ; *Search ; Search += strlen(Search) + 1) {
if (strncmp (Search, Key, n) == 0) {
if (Search[n] == '=') return Search + n + 1 ;
if (Search[n] == 0) return (char *) "1" ;
}
}
return NULL ;
}
static int kvGetInt (char * kvList, const char * Key, int Default) {
char * v = kvGet (kvList, Key) ;
int rslt = v ? ::strtol (v, NULL, 0) : Default ;
if (Knob_ReportSettings && v != NULL) {
::printf (" SyncKnob: %s %d(%d)\n", Key, rslt, Default) ;
::fflush (stdout) ;
}
return rslt ;
}
// By convention we unlink a contending thread from EntryList|cxq immediately
// after the thread acquires the lock in ::enter(). Equally, we could defer
// unlinking the thread until ::exit()-time.
void ObjectMonitor::UnlinkAfterAcquire (Thread * Self, ObjectWaiter * SelfNode)
{
assert (_owner == Self, "invariant") ;
assert (SelfNode->_thread == Self, "invariant") ;
if (SelfNode->TState == ObjectWaiter::TS_ENTER) {
// Normal case: remove Self from the DLL EntryList .
// This is a constant-time operation.
ObjectWaiter * nxt = SelfNode->_next ;
ObjectWaiter * prv = SelfNode->_prev ;
if (nxt != NULL) nxt->_prev = prv ;
if (prv != NULL) prv->_next = nxt ;
if (SelfNode == _EntryList ) _EntryList = nxt ;
assert (nxt == NULL || nxt->TState == ObjectWaiter::TS_ENTER, "invariant") ;
assert (prv == NULL || prv->TState == ObjectWaiter::TS_ENTER, "invariant") ;
TEVENT (Unlink from EntryList) ;
} else {
guarantee (SelfNode->TState == ObjectWaiter::TS_CXQ, "invariant") ;
// Inopportune interleaving -- Self is still on the cxq.
// This usually means the enqueue of self raced an exiting thread.
// Normally we'll find Self near the front of the cxq, so
// dequeueing is typically fast. If needbe we can accelerate
// this with some MCS/CHL-like bidirectional list hints and advisory
// back-links so dequeueing from the interior will normally operate
// in constant-time.
// Dequeue Self from either the head (with CAS) or from the interior
// with a linear-time scan and normal non-atomic memory operations.
// CONSIDER: if Self is on the cxq then simply drain cxq into EntryList
// and then unlink Self from EntryList. We have to drain eventually,
// so it might as well be now.
ObjectWaiter * v = _cxq ;
assert (v != NULL, "invariant") ;
if (v != SelfNode || Atomic::cmpxchg_ptr (SelfNode->_next, &_cxq, v) != v) {
// The CAS above can fail from interference IFF a "RAT" arrived.
// In that case Self must be in the interior and can no longer be
// at the head of cxq.
if (v == SelfNode) {
assert (_cxq != v, "invariant") ;
v = _cxq ; // CAS above failed - start scan at head of list
}
ObjectWaiter * p ;
ObjectWaiter * q = NULL ;
for (p = v ; p != NULL && p != SelfNode; p = p->_next) {
q = p ;
assert (p->TState == ObjectWaiter::TS_CXQ, "invariant") ;
}
assert (v != SelfNode, "invariant") ;
assert (p == SelfNode, "Node not found on cxq") ;
assert (p != _cxq, "invariant") ;
assert (q != NULL, "invariant") ;
assert (q->_next == p, "invariant") ;
q->_next = p->_next ;
}
TEVENT (Unlink from cxq) ;
}
// Diagnostic hygiene ...
SelfNode->_prev = (ObjectWaiter *) 0xBAD ;
SelfNode->_next = (ObjectWaiter *) 0xBAD ;
SelfNode->TState = ObjectWaiter::TS_RUN ;
}
// Caveat: TryLock() is not necessarily serializing if it returns failure.
// Callers must compensate as needed.
int ObjectMonitor::TryLock (Thread * Self) {
for (;;) {
void * own = _owner ;
if (own != NULL) return 0 ;
if (Atomic::cmpxchg_ptr (Self, &_owner, NULL) == NULL) {
// Either guarantee _recursions == 0 or set _recursions = 0.
assert (_recursions == 0, "invariant") ;
assert (_owner == Self, "invariant") ;
// CONSIDER: set or assert that OwnerIsThread == 1
return 1 ;
}
// The lock had been free momentarily, but we lost the race to the lock.
// Interference -- the CAS failed.
// We can either return -1 or retry.
// Retry doesn't make as much sense because the lock was just acquired.
if (true) return -1 ;
}
}
// NotRunnable() -- informed spinning
//
// Don't bother spinning if the owner is not eligible to drop the lock.
// Peek at the owner's schedctl.sc_state and Thread._thread_values and
// spin only if the owner thread is _thread_in_Java or _thread_in_vm.
// The thread must be runnable in order to drop the lock in timely fashion.
// If the _owner is not runnable then spinning will not likely be
// successful (profitable).
//
// Beware -- the thread referenced by _owner could have died
// so a simply fetch from _owner->_thread_state might trap.
// Instead, we use SafeFetchXX() to safely LD _owner->_thread_state.
// Because of the lifecycle issues the schedctl and _thread_state values
// observed by NotRunnable() might be garbage. NotRunnable must
// tolerate this and consider the observed _thread_state value
// as advisory.
//
// Beware too, that _owner is sometimes a BasicLock address and sometimes
// a thread pointer. We differentiate the two cases with OwnerIsThread.
// Alternately, we might tag the type (thread pointer vs basiclock pointer)
// with the LSB of _owner. Another option would be to probablistically probe
// the putative _owner->TypeTag value.
//
// Checking _thread_state isn't perfect. Even if the thread is
// in_java it might be blocked on a page-fault or have been preempted
// and sitting on a ready/dispatch queue. _thread state in conjunction
// with schedctl.sc_state gives us a good picture of what the
// thread is doing, however.
//
// TODO: check schedctl.sc_state.
// We'll need to use SafeFetch32() to read from the schedctl block.
// See RFE #5004247 and http://sac.sfbay.sun.com/Archives/CaseLog/arc/PSARC/2005/351/
//
// The return value from NotRunnable() is *advisory* -- the
// result is based on sampling and is not necessarily coherent.
// The caller must tolerate false-negative and false-positive errors.
// Spinning, in general, is probabilistic anyway.
int ObjectMonitor::NotRunnable (Thread * Self, Thread * ox) {
// Check either OwnerIsThread or ox->TypeTag == 2BAD.
if (!OwnerIsThread) return 0 ;
if (ox == NULL) return 0 ;
// Avoid transitive spinning ...
// Say T1 spins or blocks trying to acquire L. T1._Stalled is set to L.
// Immediately after T1 acquires L it's possible that T2, also
// spinning on L, will see L.Owner=T1 and T1._Stalled=L.
// This occurs transiently after T1 acquired L but before
// T1 managed to clear T1.Stalled. T2 does not need to abort
// its spin in this circumstance.
intptr_t BlockedOn = SafeFetchN ((intptr_t *) &ox->_Stalled, intptr_t(1)) ;
if (BlockedOn == 1) return 1 ;
if (BlockedOn != 0) {
return BlockedOn != intptr_t(this) && _owner == ox ;
}
assert (sizeof(((JavaThread *)ox)->_thread_state == sizeof(int)), "invariant") ;
int jst = SafeFetch32 ((int *) &((JavaThread *) ox)->_thread_state, -1) ; ;
// consider also: jst != _thread_in_Java -- but that's overspecific.
return jst == _thread_blocked || jst == _thread_in_native ;
}
// Adaptive spin-then-block - rational spinning
//
// Note that we spin "globally" on _owner with a classic SMP-polite TATAS
// algorithm. On high order SMP systems it would be better to start with
// a brief global spin and then revert to spinning locally. In the spirit of MCS/CLH,
// a contending thread could enqueue itself on the cxq and then spin locally
// on a thread-specific variable such as its ParkEvent._Event flag.
// That's left as an exercise for the reader. Note that global spinning is
// not problematic on Niagara, as the L2$ serves the interconnect and has both
// low latency and massive bandwidth.
//
// Broadly, we can fix the spin frequency -- that is, the % of contended lock
// acquisition attempts where we opt to spin -- at 100% and vary the spin count
// (duration) or we can fix the count at approximately the duration of
// a context switch and vary the frequency. Of course we could also
// vary both satisfying K == Frequency * Duration, where K is adaptive by monitor.
// See http://j2se.east/~dice/PERSIST/040824-AdaptiveSpinning.html.
//
// This implementation varies the duration "D", where D varies with
// the success rate of recent spin attempts. (D is capped at approximately
// length of a round-trip context switch). The success rate for recent
// spin attempts is a good predictor of the success rate of future spin
// attempts. The mechanism adapts automatically to varying critical
// section length (lock modality), system load and degree of parallelism.
// D is maintained per-monitor in _SpinDuration and is initialized
// optimistically. Spin frequency is fixed at 100%.
//
// Note that _SpinDuration is volatile, but we update it without locks
// or atomics. The code is designed so that _SpinDuration stays within
// a reasonable range even in the presence of races. The arithmetic
// operations on _SpinDuration are closed over the domain of legal values,
// so at worst a race will install and older but still legal value.
// At the very worst this introduces some apparent non-determinism.
// We might spin when we shouldn't or vice-versa, but since the spin
// count are relatively short, even in the worst case, the effect is harmless.
//
// Care must be taken that a low "D" value does not become an
// an absorbing state. Transient spinning failures -- when spinning
// is overall profitable -- should not cause the system to converge
// on low "D" values. We want spinning to be stable and predictable
// and fairly responsive to change and at the same time we don't want
// it to oscillate, become metastable, be "too" non-deterministic,
// or converge on or enter undesirable stable absorbing states.
//
// We implement a feedback-based control system -- using past behavior
// to predict future behavior. We face two issues: (a) if the
// input signal is random then the spin predictor won't provide optimal
// results, and (b) if the signal frequency is too high then the control
// system, which has some natural response lag, will "chase" the signal.
// (b) can arise from multimodal lock hold times. Transient preemption
// can also result in apparent bimodal lock hold times.
// Although sub-optimal, neither condition is particularly harmful, as
// in the worst-case we'll spin when we shouldn't or vice-versa.
// The maximum spin duration is rather short so the failure modes aren't bad.
// To be conservative, I've tuned the gain in system to bias toward
// _not spinning. Relatedly, the system can sometimes enter a mode where it
// "rings" or oscillates between spinning and not spinning. This happens
// when spinning is just on the cusp of profitability, however, so the
// situation is not dire. The state is benign -- there's no need to add
// hysteresis control to damp the transition rate between spinning and
// not spinning.
//
// - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - - -
//
// Spin-then-block strategies ...
//
// Thoughts on ways to improve spinning :
//
// * Periodically call {psr_}getloadavg() while spinning, and
// permit unbounded spinning if the load average is <
// the number of processors. Beware, however, that getloadavg()
// is exceptionally fast on solaris (about 1/10 the cost of a full
// spin cycle, but quite expensive on linux. Beware also, that
// multiple JVMs could "ring" or oscillate in a feedback loop.
// Sufficient damping would solve that problem.
//
// * We currently use spin loops with iteration counters to approximate
// spinning for some interval. Given the availability of high-precision
// time sources such as gethrtime(), %TICK, %STICK, RDTSC, etc., we should
// someday reimplement the spin loops to duration-based instead of iteration-based.
//
// * Don't spin if there are more than N = (CPUs/2) threads
// currently spinning on the monitor (or globally).
// That is, limit the number of concurrent spinners.
// We might also limit the # of spinners in the JVM, globally.
//
// * If a spinning thread observes _owner change hands it should
// abort the spin (and park immediately) or at least debit
// the spin counter by a large "penalty".
//
// * Classically, the spin count is either K*(CPUs-1) or is a
// simple constant that approximates the length of a context switch.
// We currently use a value -- computed by a special utility -- that
// approximates round-trip context switch times.
//
// * Normally schedctl_start()/_stop() is used to advise the kernel
// to avoid preempting threads that are running in short, bounded
// critical sections. We could use the schedctl hooks in an inverted
// sense -- spinners would set the nopreempt flag, but poll the preempt
// pending flag. If a spinner observed a pending preemption it'd immediately
// abort the spin and park. As such, the schedctl service acts as
// a preemption warning mechanism.
//
// * In lieu of spinning, if the system is running below saturation
// (that is, loadavg() << #cpus), we can instead suppress futile
// wakeup throttling, or even wake more than one successor at exit-time.
// The net effect is largely equivalent to spinning. In both cases,
// contending threads go ONPROC and opportunistically attempt to acquire
// the lock, decreasing lock handover latency at the expense of wasted
// cycles and context switching.
//
// * We might to spin less after we've parked as the thread will
// have less $ and TLB affinity with the processor.
// Likewise, we might spin less if we come ONPROC on a different
// processor or after a long period (>> rechose_interval).
//
// * A table-driven state machine similar to Solaris' dispadmin scheduling
// tables might be a better design. Instead of encoding information in
// _SpinDuration, _SpinFreq and _SpinClock we'd just use explicit,
// discrete states. Success or failure during a spin would drive
// state transitions, and each state node would contain a spin count.
//
// * If the processor is operating in a mode intended to conserve power
// (such as Intel's SpeedStep) or to reduce thermal output (thermal
// step-down mode) then the Java synchronization subsystem should
// forgo spinning.
//
// * The minimum spin duration should be approximately the worst-case
// store propagation latency on the platform. That is, the time
// it takes a store on CPU A to become visible on CPU B, where A and
// B are "distant".
//
// * We might want to factor a thread's priority in the spin policy.
// Threads with a higher priority might spin for slightly longer.
// Similarly, if we use back-off in the TATAS loop, lower priority
// threads might back-off longer. We don't currently use a
// thread's priority when placing it on the entry queue. We may
// want to consider doing so in future releases.
//
// * We might transiently drop a thread's scheduling priority while it spins.
// SCHED_BATCH on linux and FX scheduling class at priority=0 on Solaris
// would suffice. We could even consider letting the thread spin indefinitely at
// a depressed or "idle" priority. This brings up fairness issues, however --
// in a saturated system a thread would with a reduced priority could languish
// for extended periods on the ready queue.
//
// * While spinning try to use the otherwise wasted time to help the VM make
// progress:
//
// -- YieldTo() the owner, if the owner is OFFPROC but ready
// Done our remaining quantum directly to the ready thread.
// This helps "push" the lock owner through the critical section.
// It also tends to improve affinity/locality as the lock
// "migrates" less frequently between CPUs.
// -- Walk our own stack in anticipation of blocking. Memoize the roots.
// -- Perform strand checking for other thread. Unpark potential strandees.
// -- Help GC: trace or mark -- this would need to be a bounded unit of work.
// Unfortunately this will pollute our $ and TLBs. Recall that we
// spin to avoid context switching -- context switching has an
// immediate cost in latency, a disruptive cost to other strands on a CMT
// processor, and an amortized cost because of the D$ and TLB cache
// reload transient when the thread comes back ONPROC and repopulates
// $s and TLBs.
// -- call getloadavg() to see if the system is saturated. It'd probably
// make sense to call getloadavg() half way through the spin.
// If the system isn't at full capacity the we'd simply reset
// the spin counter to and extend the spin attempt.
// -- Doug points out that we should use the same "helping" policy
// in thread.yield().
//
// * Try MONITOR-MWAIT on systems that support those instructions.
//
// * The spin statistics that drive spin decisions & frequency are
// maintained in the objectmonitor structure so if we deflate and reinflate
// we lose spin state. In practice this is not usually a concern
// as the default spin state after inflation is aggressive (optimistic)
// and tends toward spinning. So in the worst case for a lock where
// spinning is not profitable we may spin unnecessarily for a brief
// period. But then again, if a lock is contended it'll tend not to deflate
// in the first place.
intptr_t ObjectMonitor::SpinCallbackArgument = 0 ;
int (*ObjectMonitor::SpinCallbackFunction)(intptr_t, int) = NULL ;
// Spinning: Fixed frequency (100%), vary duration
int ObjectMonitor::TrySpin_VaryDuration (Thread * Self) {
// Dumb, brutal spin. Good for comparative measurements against adaptive spinning.
int ctr = Knob_FixedSpin ;
if (ctr != 0) {
while (--ctr >= 0) {
if (TryLock (Self) > 0) return 1 ;
SpinPause () ;
}
return 0 ;
}
for (ctr = Knob_PreSpin + 1; --ctr >= 0 ; ) {
if (TryLock(Self) > 0) {
// Increase _SpinDuration ...
// Note that we don't clamp SpinDuration precisely at SpinLimit.
// Raising _SpurDuration to the poverty line is key.
int x = _SpinDuration ;
if (x < Knob_SpinLimit) {
if (x < Knob_Poverty) x = Knob_Poverty ;
_SpinDuration = x + Knob_BonusB ;
}
return 1 ;
}
SpinPause () ;
}
// Admission control - verify preconditions for spinning
//
// We always spin a little bit, just to prevent _SpinDuration == 0 from
// becoming an absorbing state. Put another way, we spin briefly to
// sample, just in case the system load, parallelism, contention, or lock
// modality changed.
//
// Consider the following alternative:
// Periodically set _SpinDuration = _SpinLimit and try a long/full
// spin attempt. "Periodically" might mean after a tally of
// the # of failed spin attempts (or iterations) reaches some threshold.
// This takes us into the realm of 1-out-of-N spinning, where we
// hold the duration constant but vary the frequency.
ctr = _SpinDuration ;
if (ctr < Knob_SpinBase) ctr = Knob_SpinBase ;
if (ctr <= 0) return 0 ;
if (Knob_SuccRestrict && _succ != NULL) return 0 ;
if (Knob_OState && NotRunnable (Self, (Thread *) _owner)) {
TEVENT (Spin abort - notrunnable [TOP]);
return 0 ;
}
int MaxSpin = Knob_MaxSpinners ;
if (MaxSpin >= 0) {
if (_Spinner > MaxSpin) {
TEVENT (Spin abort -- too many spinners) ;
return 0 ;
}
// Slighty racy, but benign ...
Adjust (&_Spinner, 1) ;
}
// We're good to spin ... spin ingress.
// CONSIDER: use Prefetch::write() to avoid RTS->RTO upgrades
// when preparing to LD...CAS _owner, etc and the CAS is likely
// to succeed.
int hits = 0 ;
int msk = 0 ;
int caspty = Knob_CASPenalty ;
int oxpty = Knob_OXPenalty ;
int sss = Knob_SpinSetSucc ;
if (sss && _succ == NULL ) _succ = Self ;
Thread * prv = NULL ;
// There are three ways to exit the following loop:
// 1. A successful spin where this thread has acquired the lock.
// 2. Spin failure with prejudice
// 3. Spin failure without prejudice
while (--ctr >= 0) {
// Periodic polling -- Check for pending GC
// Threads may spin while they're unsafe.
// We don't want spinning threads to delay the JVM from reaching
// a stop-the-world safepoint or to steal cycles from GC.
// If we detect a pending safepoint we abort in order that
// (a) this thread, if unsafe, doesn't delay the safepoint, and (b)
// this thread, if safe, doesn't steal cycles from GC.
// This is in keeping with the "no loitering in runtime" rule.
// We periodically check to see if there's a safepoint pending.
if ((ctr & 0xFF) == 0) {
if (SafepointSynchronize::do_call_back()) {
TEVENT (Spin: safepoint) ;
goto Abort ; // abrupt spin egress
}
if (Knob_UsePause & 1) SpinPause () ;
int (*scb)(intptr_t,int) = SpinCallbackFunction ;
if (hits > 50 && scb != NULL) {
int abend = (*scb)(SpinCallbackArgument, 0) ;
}
}
if (Knob_UsePause & 2) SpinPause() ;
// Exponential back-off ... Stay off the bus to reduce coherency traffic.
// This is useful on classic SMP systems, but is of less utility on
// N1-style CMT platforms.
//
// Trade-off: lock acquisition latency vs coherency bandwidth.
// Lock hold times are typically short. A histogram
// of successful spin attempts shows that we usually acquire
// the lock early in the spin. That suggests we want to
// sample _owner frequently in the early phase of the spin,
// but then back-off and sample less frequently as the spin
// progresses. The back-off makes a good citizen on SMP big
// SMP systems. Oversampling _owner can consume excessive
// coherency bandwidth. Relatedly, if we _oversample _owner we
// can inadvertently interfere with the the ST m->owner=null.
// executed by the lock owner.
if (ctr & msk) continue ;
++hits ;
if ((hits & 0xF) == 0) {
// The 0xF, above, corresponds to the exponent.
// Consider: (msk+1)|msk
msk = ((msk << 2)|3) & BackOffMask ;
}
// Probe _owner with TATAS
// If this thread observes the monitor transition or flicker
// from locked to unlocked to locked, then the odds that this
// thread will acquire the lock in this spin attempt go down
// considerably. The same argument applies if the CAS fails
// or if we observe _owner change from one non-null value to
// another non-null value. In such cases we might abort
// the spin without prejudice or apply a "penalty" to the
// spin count-down variable "ctr", reducing it by 100, say.
Thread * ox = (Thread *) _owner ;
if (ox == NULL) {
ox = (Thread *) Atomic::cmpxchg_ptr (Self, &_owner, NULL) ;
if (ox == NULL) {
// The CAS succeeded -- this thread acquired ownership
// Take care of some bookkeeping to exit spin state.
if (sss && _succ == Self) {
_succ = NULL ;
}
if (MaxSpin > 0) Adjust (&_Spinner, -1) ;
// Increase _SpinDuration :
// The spin was successful (profitable) so we tend toward
// longer spin attempts in the future.
// CONSIDER: factor "ctr" into the _SpinDuration adjustment.
// If we acquired the lock early in the spin cycle it
// makes sense to increase _SpinDuration proportionally.
// Note that we don't clamp SpinDuration precisely at SpinLimit.
int x = _SpinDuration ;
if (x < Knob_SpinLimit) {
if (x < Knob_Poverty) x = Knob_Poverty ;
_SpinDuration = x + Knob_Bonus ;
}
return 1 ;
}
// The CAS failed ... we can take any of the following actions:
// * penalize: ctr -= Knob_CASPenalty
// * exit spin with prejudice -- goto Abort;
// * exit spin without prejudice.
// * Since CAS is high-latency, retry again immediately.
prv = ox ;
TEVENT (Spin: cas failed) ;
if (caspty == -2) break ;
if (caspty == -1) goto Abort ;
ctr -= caspty ;
continue ;
}
// Did lock ownership change hands ?
if (ox != prv && prv != NULL ) {
TEVENT (spin: Owner changed)
if (oxpty == -2) break ;
if (oxpty == -1) goto Abort ;
ctr -= oxpty ;
}
prv = ox ;
// Abort the spin if the owner is not executing.
// The owner must be executing in order to drop the lock.
// Spinning while the owner is OFFPROC is idiocy.
// Consider: ctr -= RunnablePenalty ;
if (Knob_OState && NotRunnable (Self, ox)) {
TEVENT (Spin abort - notrunnable);
goto Abort ;
}
if (sss && _succ == NULL ) _succ = Self ;
}
// Spin failed with prejudice -- reduce _SpinDuration.
// TODO: Use an AIMD-like policy to adjust _SpinDuration.
// AIMD is globally stable.
TEVENT (Spin failure) ;
{
int x = _SpinDuration ;
if (x > 0) {
// Consider an AIMD scheme like: x -= (x >> 3) + 100
// This is globally sample and tends to damp the response.
x -= Knob_Penalty ;
if (x < 0) x = 0 ;
_SpinDuration = x ;
}
}
Abort:
if (MaxSpin >= 0) Adjust (&_Spinner, -1) ;
if (sss && _succ == Self) {
_succ = NULL ;
// Invariant: after setting succ=null a contending thread
// must recheck-retry _owner before parking. This usually happens
// in the normal usage of TrySpin(), but it's safest
// to make TrySpin() as foolproof as possible.
OrderAccess::fence() ;
if (TryLock(Self) > 0) return 1 ;
}
return 0 ;
}
#define TrySpin TrySpin_VaryDuration
static void DeferredInitialize () {
if (InitDone > 0) return ;
if (Atomic::cmpxchg (-1, &InitDone, 0) != 0) {
while (InitDone != 1) ;
return ;
}
// One-shot global initialization ...
// The initialization is idempotent, so we don't need locks.
// In the future consider doing this via os::init_2().
// SyncKnobs consist of <Key>=<Value> pairs in the style
// of environment variables. Start by converting ':' to NUL.
if (SyncKnobs == NULL) SyncKnobs = "" ;
size_t sz = strlen (SyncKnobs) ;
char * knobs = (char *) malloc (sz + 2) ;
if (knobs == NULL) {
vm_exit_out_of_memory (sz + 2, "Parse SyncKnobs") ;
guarantee (0, "invariant") ;
}
strcpy (knobs, SyncKnobs) ;
knobs[sz+1] = 0 ;
for (char * p = knobs ; *p ; p++) {
if (*p == ':') *p = 0 ;
}
#define SETKNOB(x) { Knob_##x = kvGetInt (knobs, #x, Knob_##x); }
SETKNOB(ReportSettings) ;
SETKNOB(Verbose) ;
SETKNOB(FixedSpin) ;
SETKNOB(SpinLimit) ;
SETKNOB(SpinBase) ;
SETKNOB(SpinBackOff);
SETKNOB(CASPenalty) ;
SETKNOB(OXPenalty) ;
SETKNOB(LogSpins) ;
SETKNOB(SpinSetSucc) ;
SETKNOB(SuccEnabled) ;
SETKNOB(SuccRestrict) ;
SETKNOB(Penalty) ;
SETKNOB(Bonus) ;
SETKNOB(BonusB) ;
SETKNOB(Poverty) ;
SETKNOB(SpinAfterFutile) ;
SETKNOB(UsePause) ;
SETKNOB(SpinEarly) ;
SETKNOB(OState) ;
SETKNOB(MaxSpinners) ;
SETKNOB(PreSpin) ;
SETKNOB(ExitPolicy) ;
SETKNOB(QMode);
SETKNOB(ResetEvent) ;
SETKNOB(MoveNotifyee) ;
SETKNOB(FastHSSEC) ;
#undef SETKNOB
if (os::is_MP()) {
BackOffMask = (1 << Knob_SpinBackOff) - 1 ;
if (Knob_ReportSettings) ::printf ("BackOffMask=%X\n", BackOffMask) ;
// CONSIDER: BackOffMask = ROUNDUP_NEXT_POWER2 (ncpus-1)
} else {
Knob_SpinLimit = 0 ;
Knob_SpinBase = 0 ;
Knob_PreSpin = 0 ;
Knob_FixedSpin = -1 ;
}
if (Knob_LogSpins == 0) {
ObjectSynchronizer::_sync_FailedSpins = NULL ;
}
free (knobs) ;
OrderAccess::fence() ;
InitDone = 1 ;
}
// Theory of operations -- Monitors lists, thread residency, etc:
//
// * A thread acquires ownership of a monitor by successfully
// CAS()ing the _owner field from null to non-null.
//
// * Invariant: A thread appears on at most one monitor list --
// cxq, EntryList or WaitSet -- at any one time.
//
// * Contending threads "push" themselves onto the cxq with CAS
// and then spin/park.
//
// * After a contending thread eventually acquires the lock it must
// dequeue itself from either the EntryList or the cxq.
//
// * The exiting thread identifies and unparks an "heir presumptive"
// tentative successor thread on the EntryList. Critically, the
// exiting thread doesn't unlink the successor thread from the EntryList.
// After having been unparked, the wakee will recontend for ownership of
// the monitor. The successor (wakee) will either acquire the lock or
// re-park itself.
//
// Succession is provided for by a policy of competitive handoff.
// The exiting thread does _not_ grant or pass ownership to the
// successor thread. (This is also referred to as "handoff" succession").
// Instead the exiting thread releases ownership and possibly wakes
// a successor, so the successor can (re)compete for ownership of the lock.
// If the EntryList is empty but the cxq is populated the exiting
// thread will drain the cxq into the EntryList. It does so by
// by detaching the cxq (installing null with CAS) and folding
// the threads from the cxq into the EntryList. The EntryList is
// doubly linked, while the cxq is singly linked because of the
// CAS-based "push" used to enqueue recently arrived threads (RATs).
//
// * Concurrency invariants:
//
// -- only the monitor owner may access or mutate the EntryList.
// The mutex property of the monitor itself protects the EntryList
// from concurrent interference.
// -- Only the monitor owner may detach the cxq.
//
// * The monitor entry list operations avoid locks, but strictly speaking
// they're not lock-free. Enter is lock-free, exit is not.
// See http://j2se.east/~dice/PERSIST/040825-LockFreeQueues.html
//
// * The cxq can have multiple concurrent "pushers" but only one concurrent
// detaching thread. This mechanism is immune from the ABA corruption.
// More precisely, the CAS-based "push" onto cxq is ABA-oblivious.
//
// * Taken together, the cxq and the EntryList constitute or form a
// single logical queue of threads stalled trying to acquire the lock.
// We use two distinct lists to improve the odds of a constant-time
// dequeue operation after acquisition (in the ::enter() epilog) and
// to reduce heat on the list ends. (c.f. Michael Scott's "2Q" algorithm).
// A key desideratum is to minimize queue & monitor metadata manipulation
// that occurs while holding the monitor lock -- that is, we want to
// minimize monitor lock holds times. Note that even a small amount of
// fixed spinning will greatly reduce the # of enqueue-dequeue operations
// on EntryList|cxq. That is, spinning relieves contention on the "inner"
// locks and monitor metadata.
//
// Cxq points to the the set of Recently Arrived Threads attempting entry.
// Because we push threads onto _cxq with CAS, the RATs must take the form of
// a singly-linked LIFO. We drain _cxq into EntryList at unlock-time when
// the unlocking thread notices that EntryList is null but _cxq is != null.
//
// The EntryList is ordered by the prevailing queue discipline and
// can be organized in any convenient fashion, such as a doubly-linked list or
// a circular doubly-linked list. Critically, we want insert and delete operations
// to operate in constant-time. If we need a priority queue then something akin
// to Solaris' sleepq would work nicely. Viz.,
// http://agg.eng/ws/on10_nightly/source/usr/src/uts/common/os/sleepq.c.
// Queue discipline is enforced at ::exit() time, when the unlocking thread
// drains the cxq into the EntryList, and orders or reorders the threads on the
// EntryList accordingly.
//
// Barring "lock barging", this mechanism provides fair cyclic ordering,
// somewhat similar to an elevator-scan.
//
// * The monitor synchronization subsystem avoids the use of native
// synchronization primitives except for the narrow platform-specific
// park-unpark abstraction. See the comments in os_solaris.cpp regarding
// the semantics of park-unpark. Put another way, this monitor implementation
// depends only on atomic operations and park-unpark. The monitor subsystem
// manages all RUNNING->BLOCKED and BLOCKED->READY transitions while the
// underlying OS manages the READY<->RUN transitions.
//
// * Waiting threads reside on the WaitSet list -- wait() puts
// the caller onto the WaitSet.
//
// * notify() or notifyAll() simply transfers threads from the WaitSet to
// either the EntryList or cxq. Subsequent exit() operations will
// unpark the notifyee. Unparking a notifee in notify() is inefficient -
// it's likely the notifyee would simply impale itself on the lock held
// by the notifier.
//
// * An interesting alternative is to encode cxq as (List,LockByte) where
// the LockByte is 0 iff the monitor is owned. _owner is simply an auxiliary
// variable, like _recursions, in the scheme. The threads or Events that form
// the list would have to be aligned in 256-byte addresses. A thread would
// try to acquire the lock or enqueue itself with CAS, but exiting threads
// could use a 1-0 protocol and simply STB to set the LockByte to 0.
// Note that is is *not* word-tearing, but it does presume that full-word
// CAS operations are coherent with intermix with STB operations. That's true
// on most common processors.
//
// * See also http://blogs.sun.com/dave
void ATTR ObjectMonitor::EnterI (TRAPS) {
Thread * Self = THREAD ;
assert (Self->is_Java_thread(), "invariant") ;
assert (((JavaThread *) Self)->thread_state() == _thread_blocked , "invariant") ;
// Try the lock - TATAS
if (TryLock (Self) > 0) {
assert (_succ != Self , "invariant") ;
assert (_owner == Self , "invariant") ;
assert (_Responsible != Self , "invariant") ;
return ;
}
DeferredInitialize () ;
// We try one round of spinning *before* enqueueing Self.
//
// If the _owner is ready but OFFPROC we could use a YieldTo()
// operation to donate the remainder of this thread's quantum
// to the owner. This has subtle but beneficial affinity
// effects.
if (TrySpin (Self) > 0) {
assert (_owner == Self , "invariant") ;
assert (_succ != Self , "invariant") ;
assert (_Responsible != Self , "invariant") ;
return ;
}
// The Spin failed -- Enqueue and park the thread ...
assert (_succ != Self , "invariant") ;
assert (_owner != Self , "invariant") ;
assert (_Responsible != Self , "invariant") ;
// Enqueue "Self" on ObjectMonitor's _cxq.
//
// Node acts as a proxy for Self.
// As an aside, if were to ever rewrite the synchronization code mostly
// in Java, WaitNodes, ObjectMonitors, and Events would become 1st-class
// Java objects. This would avoid awkward lifecycle and liveness issues,
// as well as eliminate a subset of ABA issues.
// TODO: eliminate ObjectWaiter and enqueue either Threads or Events.
//
ObjectWaiter node(Self) ;
Self->_ParkEvent->reset() ;
node._prev = (ObjectWaiter *) 0xBAD ;
node.TState = ObjectWaiter::TS_CXQ ;
// Push "Self" onto the front of the _cxq.
// Once on cxq/EntryList, Self stays on-queue until it acquires the lock.
// Note that spinning tends to reduce the rate at which threads
// enqueue and dequeue on EntryList|cxq.
ObjectWaiter * nxt ;
for (;;) {
node._next = nxt = _cxq ;
if (Atomic::cmpxchg_ptr (&node, &_cxq, nxt) == nxt) break ;
// Interference - the CAS failed because _cxq changed. Just retry.
// As an optional optimization we retry the lock.
if (TryLock (Self) > 0) {
assert (_succ != Self , "invariant") ;
assert (_owner == Self , "invariant") ;
assert (_Responsible != Self , "invariant") ;
return ;
}
}
// Check for cxq|EntryList edge transition to non-null. This indicates
// the onset of contention. While contention persists exiting threads
// will use a ST:MEMBAR:LD 1-1 exit protocol. When contention abates exit
// operations revert to the faster 1-0 mode. This enter operation may interleave
// (race) a concurrent 1-0 exit operation, resulting in stranding, so we
// arrange for one of the contending thread to use a timed park() operations
// to detect and recover from the race. (Stranding is form of progress failure
// where the monitor is unlocked but all the contending threads remain parked).
// That is, at least one of the contended threads will periodically poll _owner.
// One of the contending threads will become the designated "Responsible" thread.
// The Responsible thread uses a timed park instead of a normal indefinite park
// operation -- it periodically wakes and checks for and recovers from potential
// strandings admitted by 1-0 exit operations. We need at most one Responsible
// thread per-monitor at any given moment. Only threads on cxq|EntryList may
// be responsible for a monitor.
//
// Currently, one of the contended threads takes on the added role of "Responsible".
// A viable alternative would be to use a dedicated "stranding checker" thread
// that periodically iterated over all the threads (or active monitors) and unparked
// successors where there was risk of stranding. This would help eliminate the
// timer scalability issues we see on some platforms as we'd only have one thread
// -- the checker -- parked on a timer.
if ((SyncFlags & 16) == 0 && nxt == NULL && _EntryList == NULL) {
// Try to assume the role of responsible thread for the monitor.
// CONSIDER: ST vs CAS vs { if (Responsible==null) Responsible=Self }
Atomic::cmpxchg_ptr (Self, &_Responsible, NULL) ;
}
// The lock have been released while this thread was occupied queueing
// itself onto _cxq. To close the race and avoid "stranding" and
// progress-liveness failure we must resample-retry _owner before parking.
// Note the Dekker/Lamport duality: ST cxq; MEMBAR; LD Owner.
// In this case the ST-MEMBAR is accomplished with CAS().
//
// TODO: Defer all thread state transitions until park-time.
// Since state transitions are heavy and inefficient we'd like
// to defer the state transitions until absolutely necessary,
// and in doing so avoid some transitions ...
TEVENT (Inflated enter - Contention) ;
int nWakeups = 0 ;
int RecheckInterval = 1 ;
for (;;) {
if (TryLock (Self) > 0) break ;
assert (_owner != Self, "invariant") ;
if ((SyncFlags & 2) && _Responsible == NULL) {
Atomic::cmpxchg_ptr (Self, &_Responsible, NULL) ;
}
// park self
if (_Responsible == Self || (SyncFlags & 1)) {
TEVENT (Inflated enter - park TIMED) ;
Self->_ParkEvent->park ((jlong) RecheckInterval) ;
// Increase the RecheckInterval, but clamp the value.
RecheckInterval *= 8 ;
if (RecheckInterval > 1000) RecheckInterval = 1000 ;
} else {
TEVENT (Inflated enter - park UNTIMED) ;
Self->_ParkEvent->park() ;
}
if (TryLock(Self) > 0) break ;
// The lock is still contested.
// Keep a tally of the # of futile wakeups.
// Note that the counter is not protected by a lock or updated by atomics.
// That is by design - we trade "lossy" counters which are exposed to
// races during updates for a lower probe effect.
TEVENT (Inflated enter - Futile wakeup) ;
if (ObjectSynchronizer::_sync_FutileWakeups != NULL) {
ObjectSynchronizer::_sync_FutileWakeups->inc() ;
}
++ nWakeups ;
// Assuming this is not a spurious wakeup we'll normally find _succ == Self.
// We can defer clearing _succ until after the spin completes
// TrySpin() must tolerate being called with _succ == Self.
// Try yet another round of adaptive spinning.
if ((Knob_SpinAfterFutile & 1) && TrySpin (Self) > 0) break ;
// We can find that we were unpark()ed and redesignated _succ while
// we were spinning. That's harmless. If we iterate and call park(),
// park() will consume the event and return immediately and we'll
// just spin again. This pattern can repeat, leaving _succ to simply
// spin on a CPU. Enable Knob_ResetEvent to clear pending unparks().
// Alternately, we can sample fired() here, and if set, forgo spinning
// in the next iteration.
if ((Knob_ResetEvent & 1) && Self->_ParkEvent->fired()) {
Self->_ParkEvent->reset() ;
OrderAccess::fence() ;
}
if (_succ == Self) _succ = NULL ;
// Invariant: after clearing _succ a thread *must* retry _owner before parking.
OrderAccess::fence() ;
}
// Egress :
// Self has acquired the lock -- Unlink Self from the cxq or EntryList.
// Normally we'll find Self on the EntryList .
// From the perspective of the lock owner (this thread), the
// EntryList is stable and cxq is prepend-only.
// The head of cxq is volatile but the interior is stable.
// In addition, Self.TState is stable.
assert (_owner == Self , "invariant") ;
assert (object() != NULL , "invariant") ;
// I'd like to write:
// guarantee (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ;
// but as we're at a safepoint that's not safe.
UnlinkAfterAcquire (Self, &node) ;
if (_succ == Self) _succ = NULL ;
assert (_succ != Self, "invariant") ;
if (_Responsible == Self) {
_Responsible = NULL ;
// Dekker pivot-point.
// Consider OrderAccess::storeload() here
// We may leave threads on cxq|EntryList without a designated
// "Responsible" thread. This is benign. When this thread subsequently
// exits the monitor it can "see" such preexisting "old" threads --
// threads that arrived on the cxq|EntryList before the fence, above --
// by LDing cxq|EntryList. Newly arrived threads -- that is, threads
// that arrive on cxq after the ST:MEMBAR, above -- will set Responsible
// non-null and elect a new "Responsible" timer thread.
//
// This thread executes:
// ST Responsible=null; MEMBAR (in enter epilog - here)
// LD cxq|EntryList (in subsequent exit)
//
// Entering threads in the slow/contended path execute:
// ST cxq=nonnull; MEMBAR; LD Responsible (in enter prolog)
// The (ST cxq; MEMBAR) is accomplished with CAS().
//
// The MEMBAR, above, prevents the LD of cxq|EntryList in the subsequent
// exit operation from floating above the ST Responsible=null.
//
// In *practice* however, EnterI() is always followed by some atomic
// operation such as the decrement of _count in ::enter(). Those atomics
// obviate the need for the explicit MEMBAR, above.
}
// We've acquired ownership with CAS().
// CAS is serializing -- it has MEMBAR/FENCE-equivalent semantics.
// But since the CAS() this thread may have also stored into _succ,
// EntryList, cxq or Responsible. These meta-data updates must be
// visible __before this thread subsequently drops the lock.
// Consider what could occur if we didn't enforce this constraint --
// STs to monitor meta-data and user-data could reorder with (become
// visible after) the ST in exit that drops ownership of the lock.
// Some other thread could then acquire the lock, but observe inconsistent
// or old monitor meta-data and heap data. That violates the JMM.
// To that end, the 1-0 exit() operation must have at least STST|LDST
// "release" barrier semantics. Specifically, there must be at least a
// STST|LDST barrier in exit() before the ST of null into _owner that drops
// the lock. The barrier ensures that changes to monitor meta-data and data
// protected by the lock will be visible before we release the lock, and
// therefore before some other thread (CPU) has a chance to acquire the lock.
// See also: http://gee.cs.oswego.edu/dl/jmm/cookbook.html.
//
// Critically, any prior STs to _succ or EntryList must be visible before
// the ST of null into _owner in the *subsequent* (following) corresponding
// monitorexit. Recall too, that in 1-0 mode monitorexit does not necessarily
// execute a serializing instruction.
if (SyncFlags & 8) {
OrderAccess::fence() ;
}
return ;
}
// ExitSuspendEquivalent:
// A faster alternate to handle_special_suspend_equivalent_condition()
//
// handle_special_suspend_equivalent_condition() unconditionally
// acquires the SR_lock. On some platforms uncontended MutexLocker()
// operations have high latency. Note that in ::enter() we call HSSEC
// while holding the monitor, so we effectively lengthen the critical sections.
//
// There are a number of possible solutions:
//
// A. To ameliorate the problem we might also defer state transitions
// to as late as possible -- just prior to parking.
// Given that, we'd call HSSEC after having returned from park(),
// but before attempting to acquire the monitor. This is only a
// partial solution. It avoids calling HSSEC while holding the
// monitor (good), but it still increases successor reacquisition latency --
// the interval between unparking a successor and the time the successor
// resumes and retries the lock. See ReenterI(), which defers state transitions.
// If we use this technique we can also avoid EnterI()-exit() loop
// in ::enter() where we iteratively drop the lock and then attempt
// to reacquire it after suspending.
//
// B. In the future we might fold all the suspend bits into a
// composite per-thread suspend flag and then update it with CAS().
// Alternately, a Dekker-like mechanism with multiple variables
// would suffice:
// ST Self->_suspend_equivalent = false
// MEMBAR
// LD Self_>_suspend_flags
//
bool ObjectMonitor::ExitSuspendEquivalent (JavaThread * jSelf) {
int Mode = Knob_FastHSSEC ;
if (Mode && !jSelf->is_external_suspend()) {
assert (jSelf->is_suspend_equivalent(), "invariant") ;
jSelf->clear_suspend_equivalent() ;
if (2 == Mode) OrderAccess::storeload() ;
if (!jSelf->is_external_suspend()) return false ;
// We raced a suspension -- fall thru into the slow path
TEVENT (ExitSuspendEquivalent - raced) ;
jSelf->set_suspend_equivalent() ;
}
return jSelf->handle_special_suspend_equivalent_condition() ;
}
// ReenterI() is a specialized inline form of the latter half of the
// contended slow-path from EnterI(). We use ReenterI() only for
// monitor reentry in wait().
//
// In the future we should reconcile EnterI() and ReenterI(), adding
// Knob_Reset and Knob_SpinAfterFutile support and restructuring the
// loop accordingly.
void ATTR ObjectMonitor::ReenterI (Thread * Self, ObjectWaiter * SelfNode) {
assert (Self != NULL , "invariant") ;
assert (SelfNode != NULL , "invariant") ;
assert (SelfNode->_thread == Self , "invariant") ;
assert (_waiters > 0 , "invariant") ;
assert (((oop)(object()))->mark() == markOopDesc::encode(this) , "invariant") ;
assert (((JavaThread *)Self)->thread_state() != _thread_blocked, "invariant") ;
JavaThread * jt = (JavaThread *) Self ;
int nWakeups = 0 ;
for (;;) {
ObjectWaiter::TStates v = SelfNode->TState ;
guarantee (v == ObjectWaiter::TS_ENTER || v == ObjectWaiter::TS_CXQ, "invariant") ;
assert (_owner != Self, "invariant") ;
if (TryLock (Self) > 0) break ;
if (TrySpin (Self) > 0) break ;
TEVENT (Wait Reentry - parking) ;
// State transition wrappers around park() ...
// ReenterI() wisely defers state transitions until
// it's clear we must park the thread.
{
OSThreadContendState osts(Self->osthread());
ThreadBlockInVM tbivm(jt);
// cleared by handle_special_suspend_equivalent_condition()
// or java_suspend_self()
jt->set_suspend_equivalent();
if (SyncFlags & 1) {
Self->_ParkEvent->park ((jlong)1000) ;
} else {
Self->_ParkEvent->park () ;
}
// were we externally suspended while we were waiting?
for (;;) {
if (!ExitSuspendEquivalent (jt)) break ;
if (_succ == Self) { _succ = NULL; OrderAccess::fence(); }
jt->java_suspend_self();
jt->set_suspend_equivalent();
}
}
// Try again, but just so we distinguish between futile wakeups and
// successful wakeups. The following test isn't algorithmically
// necessary, but it helps us maintain sensible statistics.
if (TryLock(Self) > 0) break ;
// The lock is still contested.
// Keep a tally of the # of futile wakeups.
// Note that the counter is not protected by a lock or updated by atomics.
// That is by design - we trade "lossy" counters which are exposed to
// races during updates for a lower probe effect.
TEVENT (Wait Reentry - futile wakeup) ;
++ nWakeups ;
// Assuming this is not a spurious wakeup we'll normally
// find that _succ == Self.
if (_succ == Self) _succ = NULL ;
// Invariant: after clearing _succ a contending thread
// *must* retry _owner before parking.
OrderAccess::fence() ;
if (ObjectSynchronizer::_sync_FutileWakeups != NULL) {
ObjectSynchronizer::_sync_FutileWakeups->inc() ;
}
}
// Self has acquired the lock -- Unlink Self from the cxq or EntryList .
// Normally we'll find Self on the EntryList.
// Unlinking from the EntryList is constant-time and atomic-free.
// From the perspective of the lock owner (this thread), the
// EntryList is stable and cxq is prepend-only.
// The head of cxq is volatile but the interior is stable.
// In addition, Self.TState is stable.
assert (_owner == Self, "invariant") ;
assert (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ;
UnlinkAfterAcquire (Self, SelfNode) ;
if (_succ == Self) _succ = NULL ;
assert (_succ != Self, "invariant") ;
SelfNode->TState = ObjectWaiter::TS_RUN ;
OrderAccess::fence() ; // see comments at the end of EnterI()
}
bool ObjectMonitor::try_enter(Thread* THREAD) {
if (THREAD != _owner) {
if (THREAD->is_lock_owned ((address)_owner)) {
assert(_recursions == 0, "internal state error");
_owner = THREAD ;
_recursions = 1 ;
OwnerIsThread = 1 ;
return true;
}
if (Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) != NULL) {
return false;
}
return true;
} else {
_recursions++;
return true;
}
}
void ATTR ObjectMonitor::enter(TRAPS) {
// The following code is ordered to check the most common cases first
// and to reduce RTS->RTO cache line upgrades on SPARC and IA32 processors.
Thread * const Self = THREAD ;
void * cur ;
cur = Atomic::cmpxchg_ptr (Self, &_owner, NULL) ;
if (cur == NULL) {
// Either ASSERT _recursions == 0 or explicitly set _recursions = 0.
assert (_recursions == 0 , "invariant") ;
assert (_owner == Self, "invariant") ;
// CONSIDER: set or assert OwnerIsThread == 1
return ;
}
if (cur == Self) {
// TODO-FIXME: check for integer overflow! BUGID 6557169.
_recursions ++ ;
return ;
}
if (Self->is_lock_owned ((address)cur)) {
assert (_recursions == 0, "internal state error");
_recursions = 1 ;
// Commute owner from a thread-specific on-stack BasicLockObject address to
// a full-fledged "Thread *".
_owner = Self ;
OwnerIsThread = 1 ;
return ;
}
// We've encountered genuine contention.
assert (Self->_Stalled == 0, "invariant") ;
Self->_Stalled = intptr_t(this) ;
// Try one round of spinning *before* enqueueing Self
// and before going through the awkward and expensive state
// transitions. The following spin is strictly optional ...
// Note that if we acquire the monitor from an initial spin
// we forgo posting JVMTI events and firing DTRACE probes.
if (Knob_SpinEarly && TrySpin (Self) > 0) {
assert (_owner == Self , "invariant") ;
assert (_recursions == 0 , "invariant") ;
assert (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ;
Self->_Stalled = 0 ;
return ;
}
assert (_owner != Self , "invariant") ;
assert (_succ != Self , "invariant") ;
assert (Self->is_Java_thread() , "invariant") ;
JavaThread * jt = (JavaThread *) Self ;
assert (!SafepointSynchronize::is_at_safepoint(), "invariant") ;
assert (jt->thread_state() != _thread_blocked , "invariant") ;
assert (this->object() != NULL , "invariant") ;
assert (_count >= 0, "invariant") ;
// Prevent deflation at STW-time. See deflate_idle_monitors() and is_busy().
// Ensure the object-monitor relationship remains stable while there's contention.
Atomic::inc_ptr(&_count);
{ // Change java thread status to indicate blocked on monitor enter.
JavaThreadBlockedOnMonitorEnterState jtbmes(jt, this);
DTRACE_MONITOR_PROBE(contended__enter, this, object(), jt);
if (JvmtiExport::should_post_monitor_contended_enter()) {
JvmtiExport::post_monitor_contended_enter(jt, this);
}
OSThreadContendState osts(Self->osthread());
ThreadBlockInVM tbivm(jt);
Self->set_current_pending_monitor(this);
// TODO-FIXME: change the following for(;;) loop to straight-line code.
for (;;) {
jt->set_suspend_equivalent();
// cleared by handle_special_suspend_equivalent_condition()
// or java_suspend_self()
EnterI (THREAD) ;
if (!ExitSuspendEquivalent(jt)) break ;
//
// We have acquired the contended monitor, but while we were
// waiting another thread suspended us. We don't want to enter
// the monitor while suspended because that would surprise the
// thread that suspended us.
//
_recursions = 0 ;
_succ = NULL ;
exit (Self) ;
jt->java_suspend_self();
}
Self->set_current_pending_monitor(NULL);
}
Atomic::dec_ptr(&_count);
assert (_count >= 0, "invariant") ;
Self->_Stalled = 0 ;
// Must either set _recursions = 0 or ASSERT _recursions == 0.
assert (_recursions == 0 , "invariant") ;
assert (_owner == Self , "invariant") ;
assert (_succ != Self , "invariant") ;
assert (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ;
// The thread -- now the owner -- is back in vm mode.
// Report the glorious news via TI,DTrace and jvmstat.
// The probe effect is non-trivial. All the reportage occurs
// while we hold the monitor, increasing the length of the critical
// section. Amdahl's parallel speedup law comes vividly into play.
//
// Another option might be to aggregate the events (thread local or
// per-monitor aggregation) and defer reporting until a more opportune
// time -- such as next time some thread encounters contention but has
// yet to acquire the lock. While spinning that thread could
// spinning we could increment JVMStat counters, etc.
DTRACE_MONITOR_PROBE(contended__entered, this, object(), jt);
if (JvmtiExport::should_post_monitor_contended_entered()) {
JvmtiExport::post_monitor_contended_entered(jt, this);
}
if (ObjectSynchronizer::_sync_ContendedLockAttempts != NULL) {
ObjectSynchronizer::_sync_ContendedLockAttempts->inc() ;
}
}
void ObjectMonitor::ExitEpilog (Thread * Self, ObjectWaiter * Wakee) {
assert (_owner == Self, "invariant") ;
// Exit protocol:
// 1. ST _succ = wakee
// 2. membar #loadstore|#storestore;
// 2. ST _owner = NULL
// 3. unpark(wakee)
_succ = Knob_SuccEnabled ? Wakee->_thread : NULL ;
ParkEvent * Trigger = Wakee->_event ;
// Hygiene -- once we've set _owner = NULL we can't safely dereference Wakee again.
// The thread associated with Wakee may have grabbed the lock and "Wakee" may be
// out-of-scope (non-extant).
Wakee = NULL ;
// Drop the lock
OrderAccess::release_store_ptr (&_owner, NULL) ;
OrderAccess::fence() ; // ST _owner vs LD in unpark()
// TODO-FIXME:
// If there's a safepoint pending the best policy would be to
// get _this thread to a safepoint and only wake the successor
// after the safepoint completed. monitorexit uses a "leaf"
// state transition, however, so this thread can't become
// safe at this point in time. (Its stack isn't walkable).
// The next best thing is to defer waking the successor by
// adding to a list of thread to be unparked after at the
// end of the forthcoming STW).
if (SafepointSynchronize::do_call_back()) {
TEVENT (unpark before SAFEPOINT) ;
}
// Possible optimizations ...
//
// * Consider: set Wakee->UnparkTime = timeNow()
// When the thread wakes up it'll compute (timeNow() - Self->UnparkTime()).
// By measuring recent ONPROC latency we can approximate the
// system load. In turn, we can feed that information back
// into the spinning & succession policies.
// (ONPROC latency correlates strongly with load).
//
// * Pull affinity:
// If the wakee is cold then transiently setting it's affinity
// to the current CPU is a good idea.
// See http://j2se.east/~dice/PERSIST/050624-PullAffinity.txt
DTRACE_MONITOR_PROBE(contended__exit, this, object(), Self);
Trigger->unpark() ;
// Maintain stats and report events to JVMTI
if (ObjectSynchronizer::_sync_Parks != NULL) {
ObjectSynchronizer::_sync_Parks->inc() ;
}
}
// exit()
// ~~~~~~
// Note that the collector can't reclaim the objectMonitor or deflate
// the object out from underneath the thread calling ::exit() as the
// thread calling ::exit() never transitions to a stable state.
// This inhibits GC, which in turn inhibits asynchronous (and
// inopportune) reclamation of "this".
//
// We'd like to assert that: (THREAD->thread_state() != _thread_blocked) ;
// There's one exception to the claim above, however. EnterI() can call
// exit() to drop a lock if the acquirer has been externally suspended.
// In that case exit() is called with _thread_state as _thread_blocked,
// but the monitor's _count field is > 0, which inhibits reclamation.
//
// 1-0 exit
// ~~~~~~~~
// ::exit() uses a canonical 1-1 idiom with a MEMBAR although some of
// the fast-path operators have been optimized so the common ::exit()
// operation is 1-0. See i486.ad fast_unlock(), for instance.
// The code emitted by fast_unlock() elides the usual MEMBAR. This
// greatly improves latency -- MEMBAR and CAS having considerable local
// latency on modern processors -- but at the cost of "stranding". Absent the
// MEMBAR, a thread in fast_unlock() can race a thread in the slow
// ::enter() path, resulting in the entering thread being stranding
// and a progress-liveness failure. Stranding is extremely rare.
// We use timers (timed park operations) & periodic polling to detect
// and recover from stranding. Potentially stranded threads periodically
// wake up and poll the lock. See the usage of the _Responsible variable.
//
// The CAS() in enter provides for safety and exclusion, while the CAS or
// MEMBAR in exit provides for progress and avoids stranding. 1-0 locking
// eliminates the CAS/MEMBAR from the exist path, but it admits stranding.
// We detect and recover from stranding with timers.
//
// If a thread transiently strands it'll park until (a) another
// thread acquires the lock and then drops the lock, at which time the
// exiting thread will notice and unpark the stranded thread, or, (b)
// the timer expires. If the lock is high traffic then the stranding latency
// will be low due to (a). If the lock is low traffic then the odds of
// stranding are lower, although the worst-case stranding latency
// is longer. Critically, we don't want to put excessive load in the
// platform's timer subsystem. We want to minimize both the timer injection
// rate (timers created/sec) as well as the number of timers active at
// any one time. (more precisely, we want to minimize timer-seconds, which is
// the integral of the # of active timers at any instant over time).
// Both impinge on OS scalability. Given that, at most one thread parked on
// a monitor will use a timer.
void ATTR ObjectMonitor::exit(TRAPS) {
Thread * Self = THREAD ;
if (THREAD != _owner) {
if (THREAD->is_lock_owned((address) _owner)) {
// Transmute _owner from a BasicLock pointer to a Thread address.
// We don't need to hold _mutex for this transition.
// Non-null to Non-null is safe as long as all readers can
// tolerate either flavor.
assert (_recursions == 0, "invariant") ;
_owner = THREAD ;
_recursions = 0 ;
OwnerIsThread = 1 ;
} else {
// NOTE: we need to handle unbalanced monitor enter/exit
// in native code by throwing an exception.
// TODO: Throw an IllegalMonitorStateException ?
TEVENT (Exit - Throw IMSX) ;
assert(false, "Non-balanced monitor enter/exit!");
if (false) {
THROW(vmSymbols::java_lang_IllegalMonitorStateException());
}
return;
}
}
if (_recursions != 0) {
_recursions--; // this is simple recursive enter
TEVENT (Inflated exit - recursive) ;
return ;
}
// Invariant: after setting Responsible=null an thread must execute
// a MEMBAR or other serializing instruction before fetching EntryList|cxq.
if ((SyncFlags & 4) == 0) {
_Responsible = NULL ;
}
for (;;) {
assert (THREAD == _owner, "invariant") ;
// Fast-path monitor exit:
//
// Observe the Dekker/Lamport duality:
// A thread in ::exit() executes:
// ST Owner=null; MEMBAR; LD EntryList|cxq.
// A thread in the contended ::enter() path executes the complementary:
// ST EntryList|cxq = nonnull; MEMBAR; LD Owner.
//
// Note that there's a benign race in the exit path. We can drop the
// lock, another thread can reacquire the lock immediately, and we can
// then wake a thread unnecessarily (yet another flavor of futile wakeup).
// This is benign, and we've structured the code so the windows are short
// and the frequency of such futile wakeups is low.
//
// We could eliminate the race by encoding both the "LOCKED" state and
// the queue head in a single word. Exit would then use either CAS to
// clear the LOCKED bit/byte. This precludes the desirable 1-0 optimization,
// however.
//
// Possible fast-path ::exit() optimization:
// The current fast-path exit implementation fetches both cxq and EntryList.
// See also i486.ad fast_unlock(). Testing has shown that two LDs
// isn't measurably slower than a single LD on any platforms.
// Still, we could reduce the 2 LDs to one or zero by one of the following:
//
// - Use _count instead of cxq|EntryList
// We intend to eliminate _count, however, when we switch
// to on-the-fly deflation in ::exit() as is used in
// Metalocks and RelaxedLocks.
//
// - Establish the invariant that cxq == null implies EntryList == null.
// set cxq == EMPTY (1) to encode the state where cxq is empty
// by EntryList != null. EMPTY is a distinguished value.
// The fast-path exit() would fetch cxq but not EntryList.
//
// - Encode succ as follows:
// succ = t : Thread t is the successor -- t is ready or is spinning.
// Exiting thread does not need to wake a successor.
// succ = 0 : No successor required -> (EntryList|cxq) == null
// Exiting thread does not need to wake a successor
// succ = 1 : Successor required -> (EntryList|cxq) != null and
// logically succ == null.
// Exiting thread must wake a successor.
//
// The 1-1 fast-exit path would appear as :
// _owner = null ; membar ;
// if (_succ == 1 && CAS (&_owner, null, Self) == null) goto SlowPath
// goto FastPathDone ;
//
// and the 1-0 fast-exit path would appear as:
// if (_succ == 1) goto SlowPath
// Owner = null ;
// goto FastPathDone
//
// - Encode the LSB of _owner as 1 to indicate that exit()
// must use the slow-path and make a successor ready.
// (_owner & 1) == 0 IFF succ != null || (EntryList|cxq) == null
// (_owner & 1) == 0 IFF succ == null && (EntryList|cxq) != null (obviously)
// The 1-0 fast exit path would read:
// if (_owner != Self) goto SlowPath
// _owner = null
// goto FastPathDone
if (Knob_ExitPolicy == 0) {
// release semantics: prior loads and stores from within the critical section
// must not float (reorder) past the following store that drops the lock.
// On SPARC that requires MEMBAR #loadstore|#storestore.
// But of course in TSO #loadstore|#storestore is not required.
// I'd like to write one of the following:
// A. OrderAccess::release() ; _owner = NULL
// B. OrderAccess::loadstore(); OrderAccess::storestore(); _owner = NULL;
// Unfortunately OrderAccess::release() and OrderAccess::loadstore() both
// store into a _dummy variable. That store is not needed, but can result
// in massive wasteful coherency traffic on classic SMP systems.
// Instead, I use release_store(), which is implemented as just a simple
// ST on x64, x86 and SPARC.
OrderAccess::release_store_ptr (&_owner, NULL) ; // drop the lock
OrderAccess::storeload() ; // See if we need to wake a successor
if ((intptr_t(_EntryList)|intptr_t(_cxq)) == 0 || _succ != NULL) {
TEVENT (Inflated exit - simple egress) ;
return ;
}
TEVENT (Inflated exit - complex egress) ;
// Normally the exiting thread is responsible for ensuring succession,
// but if other successors are ready or other entering threads are spinning
// then this thread can simply store NULL into _owner and exit without
// waking a successor. The existence of spinners or ready successors
// guarantees proper succession (liveness). Responsibility passes to the
// ready or running successors. The exiting thread delegates the duty.
// More precisely, if a successor already exists this thread is absolved
// of the responsibility of waking (unparking) one.
//
// The _succ variable is critical to reducing futile wakeup frequency.
// _succ identifies the "heir presumptive" thread that has been made
// ready (unparked) but that has not yet run. We need only one such
// successor thread to guarantee progress.
// See http://www.usenix.org/events/jvm01/full_papers/dice/dice.pdf
// section 3.3 "Futile Wakeup Throttling" for details.
//
// Note that spinners in Enter() also set _succ non-null.
// In the current implementation spinners opportunistically set
// _succ so that exiting threads might avoid waking a successor.
// Another less appealing alternative would be for the exiting thread
// to drop the lock and then spin briefly to see if a spinner managed
// to acquire the lock. If so, the exiting thread could exit
// immediately without waking a successor, otherwise the exiting
// thread would need to dequeue and wake a successor.
// (Note that we'd need to make the post-drop spin short, but no
// shorter than the worst-case round-trip cache-line migration time.
// The dropped lock needs to become visible to the spinner, and then
// the acquisition of the lock by the spinner must become visible to
// the exiting thread).
//
// It appears that an heir-presumptive (successor) must be made ready.
// Only the current lock owner can manipulate the EntryList or
// drain _cxq, so we need to reacquire the lock. If we fail
// to reacquire the lock the responsibility for ensuring succession
// falls to the new owner.
//
if (Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) != NULL) {
return ;
}
TEVENT (Exit - Reacquired) ;
} else {
if ((intptr_t(_EntryList)|intptr_t(_cxq)) == 0 || _succ != NULL) {
OrderAccess::release_store_ptr (&_owner, NULL) ; // drop the lock
OrderAccess::storeload() ;
// Ratify the previously observed values.
if (_cxq == NULL || _succ != NULL) {
TEVENT (Inflated exit - simple egress) ;
return ;
}
// inopportune interleaving -- the exiting thread (this thread)
// in the fast-exit path raced an entering thread in the slow-enter
// path.
// We have two choices:
// A. Try to reacquire the lock.
// If the CAS() fails return immediately, otherwise
// we either restart/rerun the exit operation, or simply
// fall-through into the code below which wakes a successor.
// B. If the elements forming the EntryList|cxq are TSM
// we could simply unpark() the lead thread and return
// without having set _succ.
if (Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) != NULL) {
TEVENT (Inflated exit - reacquired succeeded) ;
return ;
}
TEVENT (Inflated exit - reacquired failed) ;
} else {
TEVENT (Inflated exit - complex egress) ;
}
}
guarantee (_owner == THREAD, "invariant") ;
// Select an appropriate successor ("heir presumptive") from the EntryList
// and make it ready. Generally we just wake the head of EntryList .
// There's no algorithmic constraint that we use the head - it's just
// a policy decision. Note that the thread at head of the EntryList
// remains at the head until it acquires the lock. This means we'll
// repeatedly wake the same thread until it manages to grab the lock.
// This is generally a good policy - if we're seeing lots of futile wakeups
// at least we're waking/rewaking a thread that's like to be hot or warm
// (have residual D$ and TLB affinity).
//
// "Wakeup locality" optimization:
// http://j2se.east/~dice/PERSIST/040825-WakeLocality.txt
// In the future we'll try to bias the selection mechanism
// to preferentially pick a thread that recently ran on
// a processor element that shares cache with the CPU on which
// the exiting thread is running. We need access to Solaris'
// schedctl.sc_cpu to make that work.
//
ObjectWaiter * w = NULL ;
int QMode = Knob_QMode ;
if (QMode == 2 && _cxq != NULL) {
// QMode == 2 : cxq has precedence over EntryList.
// Try to directly wake a successor from the cxq.
// If successful, the successor will need to unlink itself from cxq.
w = _cxq ;
assert (w != NULL, "invariant") ;
assert (w->TState == ObjectWaiter::TS_CXQ, "Invariant") ;
ExitEpilog (Self, w) ;
return ;
}
if (QMode == 3 && _cxq != NULL) {
// Aggressively drain cxq into EntryList at the first opportunity.
// This policy ensure that recently-run threads live at the head of EntryList.
// Drain _cxq into EntryList - bulk transfer.
// First, detach _cxq.
// The following loop is tantamount to: w = swap (&cxq, NULL)
w = _cxq ;
for (;;) {
assert (w != NULL, "Invariant") ;
ObjectWaiter * u = (ObjectWaiter *) Atomic::cmpxchg_ptr (NULL, &_cxq, w) ;
if (u == w) break ;
w = u ;
}
assert (w != NULL , "invariant") ;
ObjectWaiter * q = NULL ;
ObjectWaiter * p ;
for (p = w ; p != NULL ; p = p->_next) {
guarantee (p->TState == ObjectWaiter::TS_CXQ, "Invariant") ;
p->TState = ObjectWaiter::TS_ENTER ;
p->_prev = q ;
q = p ;
}
// Append the RATs to the EntryList
// TODO: organize EntryList as a CDLL so we can locate the tail in constant-time.
ObjectWaiter * Tail ;
for (Tail = _EntryList ; Tail != NULL && Tail->_next != NULL ; Tail = Tail->_next) ;
if (Tail == NULL) {
_EntryList = w ;
} else {
Tail->_next = w ;
w->_prev = Tail ;
}
// Fall thru into code that tries to wake a successor from EntryList
}
if (QMode == 4 && _cxq != NULL) {
// Aggressively drain cxq into EntryList at the first opportunity.
// This policy ensure that recently-run threads live at the head of EntryList.
// Drain _cxq into EntryList - bulk transfer.
// First, detach _cxq.
// The following loop is tantamount to: w = swap (&cxq, NULL)
w = _cxq ;
for (;;) {
assert (w != NULL, "Invariant") ;
ObjectWaiter * u = (ObjectWaiter *) Atomic::cmpxchg_ptr (NULL, &_cxq, w) ;
if (u == w) break ;
w = u ;
}
assert (w != NULL , "invariant") ;
ObjectWaiter * q = NULL ;
ObjectWaiter * p ;
for (p = w ; p != NULL ; p = p->_next) {
guarantee (p->TState == ObjectWaiter::TS_CXQ, "Invariant") ;
p->TState = ObjectWaiter::TS_ENTER ;
p->_prev = q ;
q = p ;
}
// Prepend the RATs to the EntryList
if (_EntryList != NULL) {
q->_next = _EntryList ;
_EntryList->_prev = q ;
}
_EntryList = w ;
// Fall thru into code that tries to wake a successor from EntryList
}
w = _EntryList ;
if (w != NULL) {
// I'd like to write: guarantee (w->_thread != Self).
// But in practice an exiting thread may find itself on the EntryList.
// Lets say thread T1 calls O.wait(). Wait() enqueues T1 on O's waitset and
// then calls exit(). Exit release the lock by setting O._owner to NULL.
// Lets say T1 then stalls. T2 acquires O and calls O.notify(). The
// notify() operation moves T1 from O's waitset to O's EntryList. T2 then
// release the lock "O". T2 resumes immediately after the ST of null into
// _owner, above. T2 notices that the EntryList is populated, so it
// reacquires the lock and then finds itself on the EntryList.
// Given all that, we have to tolerate the circumstance where "w" is
// associated with Self.
assert (w->TState == ObjectWaiter::TS_ENTER, "invariant") ;
ExitEpilog (Self, w) ;
return ;
}
// If we find that both _cxq and EntryList are null then just
// re-run the exit protocol from the top.
w = _cxq ;
if (w == NULL) continue ;
// Drain _cxq into EntryList - bulk transfer.
// First, detach _cxq.
// The following loop is tantamount to: w = swap (&cxq, NULL)
for (;;) {
assert (w != NULL, "Invariant") ;
ObjectWaiter * u = (ObjectWaiter *) Atomic::cmpxchg_ptr (NULL, &_cxq, w) ;
if (u == w) break ;
w = u ;
}
TEVENT (Inflated exit - drain cxq into EntryList) ;
assert (w != NULL , "invariant") ;
assert (_EntryList == NULL , "invariant") ;
// Convert the LIFO SLL anchored by _cxq into a DLL.
// The list reorganization step operates in O(LENGTH(w)) time.
// It's critical that this step operate quickly as
// "Self" still holds the outer-lock, restricting parallelism
// and effectively lengthening the critical section.
// Invariant: s chases t chases u.
// TODO-FIXME: consider changing EntryList from a DLL to a CDLL so
// we have faster access to the tail.
if (QMode == 1) {
// QMode == 1 : drain cxq to EntryList, reversing order
// We also reverse the order of the list.
ObjectWaiter * s = NULL ;
ObjectWaiter * t = w ;
ObjectWaiter * u = NULL ;
while (t != NULL) {
guarantee (t->TState == ObjectWaiter::TS_CXQ, "invariant") ;
t->TState = ObjectWaiter::TS_ENTER ;
u = t->_next ;
t->_prev = u ;
t->_next = s ;
s = t;
t = u ;
}
_EntryList = s ;
assert (s != NULL, "invariant") ;
} else {
// QMode == 0 or QMode == 2
_EntryList = w ;
ObjectWaiter * q = NULL ;
ObjectWaiter * p ;
for (p = w ; p != NULL ; p = p->_next) {
guarantee (p->TState == ObjectWaiter::TS_CXQ, "Invariant") ;
p->TState = ObjectWaiter::TS_ENTER ;
p->_prev = q ;
q = p ;
}
}
// In 1-0 mode we need: ST EntryList; MEMBAR #storestore; ST _owner = NULL
// The MEMBAR is satisfied by the release_store() operation in ExitEpilog().
// See if we can abdicate to a spinner instead of waking a thread.
// A primary goal of the implementation is to reduce the
// context-switch rate.
if (_succ != NULL) continue;
w = _EntryList ;
if (w != NULL) {
guarantee (w->TState == ObjectWaiter::TS_ENTER, "invariant") ;
ExitEpilog (Self, w) ;
return ;
}
}
}
// complete_exit exits a lock returning recursion count
// complete_exit/reenter operate as a wait without waiting
// complete_exit requires an inflated monitor
// The _owner field is not always the Thread addr even with an
// inflated monitor, e.g. the monitor can be inflated by a non-owning
// thread due to contention.
intptr_t ObjectMonitor::complete_exit(TRAPS) {
Thread * const Self = THREAD;
assert(Self->is_Java_thread(), "Must be Java thread!");
JavaThread *jt = (JavaThread *)THREAD;
DeferredInitialize();
if (THREAD != _owner) {
if (THREAD->is_lock_owned ((address)_owner)) {
assert(_recursions == 0, "internal state error");
_owner = THREAD ; /* Convert from basiclock addr to Thread addr */
_recursions = 0 ;
OwnerIsThread = 1 ;
}
}
guarantee(Self == _owner, "complete_exit not owner");
intptr_t save = _recursions; // record the old recursion count
_recursions = 0; // set the recursion level to be 0
exit (Self) ; // exit the monitor
guarantee (_owner != Self, "invariant");
return save;
}
// reenter() enters a lock and sets recursion count
// complete_exit/reenter operate as a wait without waiting
void ObjectMonitor::reenter(intptr_t recursions, TRAPS) {
Thread * const Self = THREAD;
assert(Self->is_Java_thread(), "Must be Java thread!");
JavaThread *jt = (JavaThread *)THREAD;
guarantee(_owner != Self, "reenter already owner");
enter (THREAD); // enter the monitor
guarantee (_recursions == 0, "reenter recursion");
_recursions = recursions;
return;
}
// Note: a subset of changes to ObjectMonitor::wait()
// will need to be replicated in complete_exit above
void ObjectMonitor::wait(jlong millis, bool interruptible, TRAPS) {
Thread * const Self = THREAD ;
assert(Self->is_Java_thread(), "Must be Java thread!");
JavaThread *jt = (JavaThread *)THREAD;
DeferredInitialize () ;
// Throw IMSX or IEX.
CHECK_OWNER();
// check for a pending interrupt
if (interruptible && Thread::is_interrupted(Self, true) && !HAS_PENDING_EXCEPTION) {
// post monitor waited event. Note that this is past-tense, we are done waiting.
if (JvmtiExport::should_post_monitor_waited()) {
// Note: 'false' parameter is passed here because the
// wait was not timed out due to thread interrupt.
JvmtiExport::post_monitor_waited(jt, this, false);
}
TEVENT (Wait - Throw IEX) ;
THROW(vmSymbols::java_lang_InterruptedException());
return ;
}
TEVENT (Wait) ;
assert (Self->_Stalled == 0, "invariant") ;
Self->_Stalled = intptr_t(this) ;
jt->set_current_waiting_monitor(this);
// create a node to be put into the queue
// Critically, after we reset() the event but prior to park(), we must check
// for a pending interrupt.
ObjectWaiter node(Self);
node.TState = ObjectWaiter::TS_WAIT ;
Self->_ParkEvent->reset() ;
OrderAccess::fence(); // ST into Event; membar ; LD interrupted-flag
// Enter the waiting queue, which is a circular doubly linked list in this case
// but it could be a priority queue or any data structure.
// _WaitSetLock protects the wait queue. Normally the wait queue is accessed only
// by the the owner of the monitor *except* in the case where park()
// returns because of a timeout of interrupt. Contention is exceptionally rare
// so we use a simple spin-lock instead of a heavier-weight blocking lock.
Thread::SpinAcquire (&_WaitSetLock, "WaitSet - add") ;
AddWaiter (&node) ;
Thread::SpinRelease (&_WaitSetLock) ;
if ((SyncFlags & 4) == 0) {
_Responsible = NULL ;
}
intptr_t save = _recursions; // record the old recursion count
_waiters++; // increment the number of waiters
_recursions = 0; // set the recursion level to be 1
exit (Self) ; // exit the monitor
guarantee (_owner != Self, "invariant") ;
// As soon as the ObjectMonitor's ownership is dropped in the exit()
// call above, another thread can enter() the ObjectMonitor, do the
// notify(), and exit() the ObjectMonitor. If the other thread's
// exit() call chooses this thread as the successor and the unpark()
// call happens to occur while this thread is posting a
// MONITOR_CONTENDED_EXIT event, then we run the risk of the event
// handler using RawMonitors and consuming the unpark().
//
// To avoid the problem, we re-post the event. This does no harm
// even if the original unpark() was not consumed because we are the
// chosen successor for this monitor.
if (node._notified != 0 && _succ == Self) {
node._event->unpark();
}
// The thread is on the WaitSet list - now park() it.
// On MP systems it's conceivable that a brief spin before we park
// could be profitable.
//
// TODO-FIXME: change the following logic to a loop of the form
// while (!timeout && !interrupted && _notified == 0) park()
int ret = OS_OK ;
int WasNotified = 0 ;
{ // State transition wrappers
OSThread* osthread = Self->osthread();
OSThreadWaitState osts(osthread, true);
{
ThreadBlockInVM tbivm(jt);
// Thread is in thread_blocked state and oop access is unsafe.
jt->set_suspend_equivalent();
if (interruptible && (Thread::is_interrupted(THREAD, false) || HAS_PENDING_EXCEPTION)) {
// Intentionally empty
} else
if (node._notified == 0) {
if (millis <= 0) {
Self->_ParkEvent->park () ;
} else {
ret = Self->_ParkEvent->park (millis) ;
}
}
// were we externally suspended while we were waiting?
if (ExitSuspendEquivalent (jt)) {
// TODO-FIXME: add -- if succ == Self then succ = null.
jt->java_suspend_self();
}
} // Exit thread safepoint: transition _thread_blocked -> _thread_in_vm
// Node may be on the WaitSet, the EntryList (or cxq), or in transition
// from the WaitSet to the EntryList.
// See if we need to remove Node from the WaitSet.
// We use double-checked locking to avoid grabbing _WaitSetLock
// if the thread is not on the wait queue.
//
// Note that we don't need a fence before the fetch of TState.
// In the worst case we'll fetch a old-stale value of TS_WAIT previously
// written by the is thread. (perhaps the fetch might even be satisfied
// by a look-aside into the processor's own store buffer, although given
// the length of the code path between the prior ST and this load that's
// highly unlikely). If the following LD fetches a stale TS_WAIT value
// then we'll acquire the lock and then re-fetch a fresh TState value.
// That is, we fail toward safety.
if (node.TState == ObjectWaiter::TS_WAIT) {
Thread::SpinAcquire (&_WaitSetLock, "WaitSet - unlink") ;
if (node.TState == ObjectWaiter::TS_WAIT) {
DequeueSpecificWaiter (&node) ; // unlink from WaitSet
assert(node._notified == 0, "invariant");
node.TState = ObjectWaiter::TS_RUN ;
}
Thread::SpinRelease (&_WaitSetLock) ;
}
// The thread is now either on off-list (TS_RUN),
// on the EntryList (TS_ENTER), or on the cxq (TS_CXQ).
// The Node's TState variable is stable from the perspective of this thread.
// No other threads will asynchronously modify TState.
guarantee (node.TState != ObjectWaiter::TS_WAIT, "invariant") ;
OrderAccess::loadload() ;
if (_succ == Self) _succ = NULL ;
WasNotified = node._notified ;
// Reentry phase -- reacquire the monitor.
// re-enter contended monitor after object.wait().
// retain OBJECT_WAIT state until re-enter successfully completes
// Thread state is thread_in_vm and oop access is again safe,
// although the raw address of the object may have changed.
// (Don't cache naked oops over safepoints, of course).
// post monitor waited event. Note that this is past-tense, we are done waiting.
if (JvmtiExport::should_post_monitor_waited()) {
JvmtiExport::post_monitor_waited(jt, this, ret == OS_TIMEOUT);
}
OrderAccess::fence() ;
assert (Self->_Stalled != 0, "invariant") ;
Self->_Stalled = 0 ;
assert (_owner != Self, "invariant") ;
ObjectWaiter::TStates v = node.TState ;
if (v == ObjectWaiter::TS_RUN) {
enter (Self) ;
} else {
guarantee (v == ObjectWaiter::TS_ENTER || v == ObjectWaiter::TS_CXQ, "invariant") ;
ReenterI (Self, &node) ;
node.wait_reenter_end(this);
}
// Self has reacquired the lock.
// Lifecycle - the node representing Self must not appear on any queues.
// Node is about to go out-of-scope, but even if it were immortal we wouldn't
// want residual elements associated with this thread left on any lists.
guarantee (node.TState == ObjectWaiter::TS_RUN, "invariant") ;
assert (_owner == Self, "invariant") ;
assert (_succ != Self , "invariant") ;
} // OSThreadWaitState()
jt->set_current_waiting_monitor(NULL);
guarantee (_recursions == 0, "invariant") ;
_recursions = save; // restore the old recursion count
_waiters--; // decrement the number of waiters
// Verify a few postconditions
assert (_owner == Self , "invariant") ;
assert (_succ != Self , "invariant") ;
assert (((oop)(object()))->mark() == markOopDesc::encode(this), "invariant") ;
if (SyncFlags & 32) {
OrderAccess::fence() ;
}
// check if the notification happened
if (!WasNotified) {
// no, it could be timeout or Thread.interrupt() or both
// check for interrupt event, otherwise it is timeout
if (interruptible && Thread::is_interrupted(Self, true) && !HAS_PENDING_EXCEPTION) {
TEVENT (Wait - throw IEX from epilog) ;
THROW(vmSymbols::java_lang_InterruptedException());
}
}
// NOTE: Spurious wake up will be consider as timeout.
// Monitor notify has precedence over thread interrupt.
}
// Consider:
// If the lock is cool (cxq == null && succ == null) and we're on an MP system
// then instead of transferring a thread from the WaitSet to the EntryList
// we might just dequeue a thread from the WaitSet and directly unpark() it.
void ObjectMonitor::notify(TRAPS) {
CHECK_OWNER();
if (_WaitSet == NULL) {
TEVENT (Empty-Notify) ;
return ;
}
DTRACE_MONITOR_PROBE(notify, this, object(), THREAD);
int Policy = Knob_MoveNotifyee ;
Thread::SpinAcquire (&_WaitSetLock, "WaitSet - notify") ;
ObjectWaiter * iterator = DequeueWaiter() ;
if (iterator != NULL) {
TEVENT (Notify1 - Transfer) ;
guarantee (iterator->TState == ObjectWaiter::TS_WAIT, "invariant") ;
guarantee (iterator->_notified == 0, "invariant") ;
// Disposition - what might we do with iterator ?
// a. add it directly to the EntryList - either tail or head.
// b. push it onto the front of the _cxq.
// For now we use (a).
if (Policy != 4) {
iterator->TState = ObjectWaiter::TS_ENTER ;
}
iterator->_notified = 1 ;
ObjectWaiter * List = _EntryList ;
if (List != NULL) {
assert (List->_prev == NULL, "invariant") ;
assert (List->TState == ObjectWaiter::TS_ENTER, "invariant") ;
assert (List != iterator, "invariant") ;
}
if (Policy == 0) { // prepend to EntryList
if (List == NULL) {
iterator->_next = iterator->_prev = NULL ;
_EntryList = iterator ;
} else {
List->_prev = iterator ;
iterator->_next = List ;
iterator->_prev = NULL ;
_EntryList = iterator ;
}
} else
if (Policy == 1) { // append to EntryList
if (List == NULL) {
iterator->_next = iterator->_prev = NULL ;
_EntryList = iterator ;
} else {
// CONSIDER: finding the tail currently requires a linear-time walk of
// the EntryList. We can make tail access constant-time by converting to
// a CDLL instead of using our current DLL.
ObjectWaiter * Tail ;
for (Tail = List ; Tail->_next != NULL ; Tail = Tail->_next) ;
assert (Tail != NULL && Tail->_next == NULL, "invariant") ;
Tail->_next = iterator ;
iterator->_prev = Tail ;
iterator->_next = NULL ;
}
} else
if (Policy == 2) { // prepend to cxq
// prepend to cxq
if (List == NULL) {
iterator->_next = iterator->_prev = NULL ;
_EntryList = iterator ;
} else {
iterator->TState = ObjectWaiter::TS_CXQ ;
for (;;) {
ObjectWaiter * Front = _cxq ;
iterator->_next = Front ;
if (Atomic::cmpxchg_ptr (iterator, &_cxq, Front) == Front) {
break ;
}
}
}
} else
if (Policy == 3) { // append to cxq
iterator->TState = ObjectWaiter::TS_CXQ ;
for (;;) {
ObjectWaiter * Tail ;
Tail = _cxq ;
if (Tail == NULL) {
iterator->_next = NULL ;
if (Atomic::cmpxchg_ptr (iterator, &_cxq, NULL) == NULL) {
break ;
}
} else {
while (Tail->_next != NULL) Tail = Tail->_next ;
Tail->_next = iterator ;
iterator->_prev = Tail ;
iterator->_next = NULL ;
break ;
}
}
} else {
ParkEvent * ev = iterator->_event ;
iterator->TState = ObjectWaiter::TS_RUN ;
OrderAccess::fence() ;
ev->unpark() ;
}
if (Policy < 4) {
iterator->wait_reenter_begin(this);
}
// _WaitSetLock protects the wait queue, not the EntryList. We could
// move the add-to-EntryList operation, above, outside the critical section
// protected by _WaitSetLock. In practice that's not useful. With the
// exception of wait() timeouts and interrupts the monitor owner
// is the only thread that grabs _WaitSetLock. There's almost no contention
// on _WaitSetLock so it's not profitable to reduce the length of the
// critical section.
}
Thread::SpinRelease (&_WaitSetLock) ;
if (iterator != NULL && ObjectSynchronizer::_sync_Notifications != NULL) {
ObjectSynchronizer::_sync_Notifications->inc() ;
}
}
void ObjectMonitor::notifyAll(TRAPS) {
CHECK_OWNER();
ObjectWaiter* iterator;
if (_WaitSet == NULL) {
TEVENT (Empty-NotifyAll) ;
return ;
}
DTRACE_MONITOR_PROBE(notifyAll, this, object(), THREAD);
int Policy = Knob_MoveNotifyee ;
int Tally = 0 ;
Thread::SpinAcquire (&_WaitSetLock, "WaitSet - notifyall") ;
for (;;) {
iterator = DequeueWaiter () ;
if (iterator == NULL) break ;
TEVENT (NotifyAll - Transfer1) ;
++Tally ;
// Disposition - what might we do with iterator ?
// a. add it directly to the EntryList - either tail or head.
// b. push it onto the front of the _cxq.
// For now we use (a).
//
// TODO-FIXME: currently notifyAll() transfers the waiters one-at-a-time from the waitset
// to the EntryList. This could be done more efficiently with a single bulk transfer,
// but in practice it's not time-critical. Beware too, that in prepend-mode we invert the
// order of the waiters. Lets say that the waitset is "ABCD" and the EntryList is "XYZ".
// After a notifyAll() in prepend mode the waitset will be empty and the EntryList will
// be "DCBAXYZ".
guarantee (iterator->TState == ObjectWaiter::TS_WAIT, "invariant") ;
guarantee (iterator->_notified == 0, "invariant") ;
iterator->_notified = 1 ;
if (Policy != 4) {
iterator->TState = ObjectWaiter::TS_ENTER ;
}
ObjectWaiter * List = _EntryList ;
if (List != NULL) {
assert (List->_prev == NULL, "invariant") ;
assert (List->TState == ObjectWaiter::TS_ENTER, "invariant") ;
assert (List != iterator, "invariant") ;
}
if (Policy == 0) { // prepend to EntryList
if (List == NULL) {
iterator->_next = iterator->_prev = NULL ;
_EntryList = iterator ;
} else {
List->_prev = iterator ;
iterator->_next = List ;
iterator->_prev = NULL ;
_EntryList = iterator ;
}
} else
if (Policy == 1) { // append to EntryList
if (List == NULL) {
iterator->_next = iterator->_prev = NULL ;
_EntryList = iterator ;
} else {
// CONSIDER: finding the tail currently requires a linear-time walk of
// the EntryList. We can make tail access constant-time by converting to
// a CDLL instead of using our current DLL.
ObjectWaiter * Tail ;
for (Tail = List ; Tail->_next != NULL ; Tail = Tail->_next) ;
assert (Tail != NULL && Tail->_next == NULL, "invariant") ;
Tail->_next = iterator ;
iterator->_prev = Tail ;
iterator->_next = NULL ;
}
} else
if (Policy == 2) { // prepend to cxq
// prepend to cxq
iterator->TState = ObjectWaiter::TS_CXQ ;
for (;;) {
ObjectWaiter * Front = _cxq ;
iterator->_next = Front ;
if (Atomic::cmpxchg_ptr (iterator, &_cxq, Front) == Front) {
break ;
}
}
} else
if (Policy == 3) { // append to cxq
iterator->TState = ObjectWaiter::TS_CXQ ;
for (;;) {
ObjectWaiter * Tail ;
Tail = _cxq ;
if (Tail == NULL) {
iterator->_next = NULL ;
if (Atomic::cmpxchg_ptr (iterator, &_cxq, NULL) == NULL) {
break ;
}
} else {
while (Tail->_next != NULL) Tail = Tail->_next ;
Tail->_next = iterator ;
iterator->_prev = Tail ;
iterator->_next = NULL ;
break ;
}
}
} else {
ParkEvent * ev = iterator->_event ;
iterator->TState = ObjectWaiter::TS_RUN ;
OrderAccess::fence() ;
ev->unpark() ;
}
if (Policy < 4) {
iterator->wait_reenter_begin(this);
}
// _WaitSetLock protects the wait queue, not the EntryList. We could
// move the add-to-EntryList operation, above, outside the critical section
// protected by _WaitSetLock. In practice that's not useful. With the
// exception of wait() timeouts and interrupts the monitor owner
// is the only thread that grabs _WaitSetLock. There's almost no contention
// on _WaitSetLock so it's not profitable to reduce the length of the
// critical section.
}
Thread::SpinRelease (&_WaitSetLock) ;
if (Tally != 0 && ObjectSynchronizer::_sync_Notifications != NULL) {
ObjectSynchronizer::_sync_Notifications->inc(Tally) ;
}
}
// check_slow() is a misnomer. It's called to simply to throw an IMSX exception.
// TODO-FIXME: remove check_slow() -- it's likely dead.
void ObjectMonitor::check_slow(TRAPS) {
TEVENT (check_slow - throw IMSX) ;
assert(THREAD != _owner && !THREAD->is_lock_owned((address) _owner), "must not be owner");
THROW_MSG(vmSymbols::java_lang_IllegalMonitorStateException(), "current thread not owner");
}
// -------------------------------------------------------------------------
// The raw monitor subsystem is entirely distinct from normal
// java-synchronization or jni-synchronization. raw monitors are not
// associated with objects. They can be implemented in any manner
// that makes sense. The original implementors decided to piggy-back
// the raw-monitor implementation on the existing Java objectMonitor mechanism.
// This flaw needs to fixed. We should reimplement raw monitors as sui-generis.
// Specifically, we should not implement raw monitors via java monitors.
// Time permitting, we should disentangle and deconvolve the two implementations
// and move the resulting raw monitor implementation over to the JVMTI directories.
// Ideally, the raw monitor implementation would be built on top of
// park-unpark and nothing else.
//
// raw monitors are used mainly by JVMTI
// The raw monitor implementation borrows the ObjectMonitor structure,
// but the operators are degenerate and extremely simple.
//
// Mixed use of a single objectMonitor instance -- as both a raw monitor
// and a normal java monitor -- is not permissible.
//
// Note that we use the single RawMonitor_lock to protect queue operations for
// _all_ raw monitors. This is a scalability impediment, but since raw monitor usage
// is deprecated and rare, this is not of concern. The RawMonitor_lock can not
// be held indefinitely. The critical sections must be short and bounded.
//
// -------------------------------------------------------------------------
int ObjectMonitor::SimpleEnter (Thread * Self) {
for (;;) {
if (Atomic::cmpxchg_ptr (Self, &_owner, NULL) == NULL) {
return OS_OK ;
}
ObjectWaiter Node (Self) ;
Self->_ParkEvent->reset() ; // strictly optional
Node.TState = ObjectWaiter::TS_ENTER ;
RawMonitor_lock->lock_without_safepoint_check() ;
Node._next = _EntryList ;
_EntryList = &Node ;
OrderAccess::fence() ;
if (_owner == NULL && Atomic::cmpxchg_ptr (Self, &_owner, NULL) == NULL) {
_EntryList = Node._next ;
RawMonitor_lock->unlock() ;
return OS_OK ;
}
RawMonitor_lock->unlock() ;
while (Node.TState == ObjectWaiter::TS_ENTER) {
Self->_ParkEvent->park() ;
}
}
}
int ObjectMonitor::SimpleExit (Thread * Self) {
guarantee (_owner == Self, "invariant") ;
OrderAccess::release_store_ptr (&_owner, NULL) ;
OrderAccess::fence() ;
if (_EntryList == NULL) return OS_OK ;
ObjectWaiter * w ;
RawMonitor_lock->lock_without_safepoint_check() ;
w = _EntryList ;
if (w != NULL) {
_EntryList = w->_next ;
}
RawMonitor_lock->unlock() ;
if (w != NULL) {
guarantee (w ->TState == ObjectWaiter::TS_ENTER, "invariant") ;
ParkEvent * ev = w->_event ;
w->TState = ObjectWaiter::TS_RUN ;
OrderAccess::fence() ;
ev->unpark() ;
}
return OS_OK ;
}
int ObjectMonitor::SimpleWait (Thread * Self, jlong millis) {
guarantee (_owner == Self , "invariant") ;
guarantee (_recursions == 0, "invariant") ;
ObjectWaiter Node (Self) ;
Node._notified = 0 ;
Node.TState = ObjectWaiter::TS_WAIT ;
RawMonitor_lock->lock_without_safepoint_check() ;
Node._next = _WaitSet ;
_WaitSet = &Node ;
RawMonitor_lock->unlock() ;
SimpleExit (Self) ;
guarantee (_owner != Self, "invariant") ;
int ret = OS_OK ;
if (millis <= 0) {
Self->_ParkEvent->park();
} else {
ret = Self->_ParkEvent->park(millis);
}
// If thread still resides on the waitset then unlink it.
// Double-checked locking -- the usage is safe in this context
// as we TState is volatile and the lock-unlock operators are
// serializing (barrier-equivalent).
if (Node.TState == ObjectWaiter::TS_WAIT) {
RawMonitor_lock->lock_without_safepoint_check() ;
if (Node.TState == ObjectWaiter::TS_WAIT) {
// Simple O(n) unlink, but performance isn't critical here.
ObjectWaiter * p ;
ObjectWaiter * q = NULL ;
for (p = _WaitSet ; p != &Node; p = p->_next) {
q = p ;
}
guarantee (p == &Node, "invariant") ;
if (q == NULL) {
guarantee (p == _WaitSet, "invariant") ;
_WaitSet = p->_next ;
} else {
guarantee (p == q->_next, "invariant") ;
q->_next = p->_next ;
}
Node.TState = ObjectWaiter::TS_RUN ;
}
RawMonitor_lock->unlock() ;
}
guarantee (Node.TState == ObjectWaiter::TS_RUN, "invariant") ;
SimpleEnter (Self) ;
guarantee (_owner == Self, "invariant") ;
guarantee (_recursions == 0, "invariant") ;
return ret ;
}
int ObjectMonitor::SimpleNotify (Thread * Self, bool All) {
guarantee (_owner == Self, "invariant") ;
if (_WaitSet == NULL) return OS_OK ;
// We have two options:
// A. Transfer the threads from the WaitSet to the EntryList
// B. Remove the thread from the WaitSet and unpark() it.
//
// We use (B), which is crude and results in lots of futile
// context switching. In particular (B) induces lots of contention.
ParkEvent * ev = NULL ; // consider using a small auto array ...
RawMonitor_lock->lock_without_safepoint_check() ;
for (;;) {
ObjectWaiter * w = _WaitSet ;
if (w == NULL) break ;
_WaitSet = w->_next ;
if (ev != NULL) { ev->unpark(); ev = NULL; }
ev = w->_event ;
OrderAccess::loadstore() ;
w->TState = ObjectWaiter::TS_RUN ;
OrderAccess::storeload();
if (!All) break ;
}
RawMonitor_lock->unlock() ;
if (ev != NULL) ev->unpark();
return OS_OK ;
}
// Any JavaThread will enter here with state _thread_blocked
int ObjectMonitor::raw_enter(TRAPS) {
TEVENT (raw_enter) ;
void * Contended ;
// don't enter raw monitor if thread is being externally suspended, it will
// surprise the suspender if a "suspended" thread can still enter monitor
JavaThread * jt = (JavaThread *)THREAD;
if (THREAD->is_Java_thread()) {
jt->SR_lock()->lock_without_safepoint_check();
while (jt->is_external_suspend()) {
jt->SR_lock()->unlock();
jt->java_suspend_self();
jt->SR_lock()->lock_without_safepoint_check();
}
// guarded by SR_lock to avoid racing with new external suspend requests.
Contended = Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) ;
jt->SR_lock()->unlock();
} else {
Contended = Atomic::cmpxchg_ptr (THREAD, &_owner, NULL) ;
}
if (Contended == THREAD) {
_recursions ++ ;
return OM_OK ;
}
if (Contended == NULL) {
guarantee (_owner == THREAD, "invariant") ;
guarantee (_recursions == 0, "invariant") ;
return OM_OK ;
}
THREAD->set_current_pending_monitor(this);
if (!THREAD->is_Java_thread()) {
// No other non-Java threads besides VM thread would acquire
// a raw monitor.
assert(THREAD->is_VM_thread(), "must be VM thread");
SimpleEnter (THREAD) ;
} else {
guarantee (jt->thread_state() == _thread_blocked, "invariant") ;
for (;;) {
jt->set_suspend_equivalent();
// cleared by handle_special_suspend_equivalent_condition() or
// java_suspend_self()
SimpleEnter (THREAD) ;
// were we externally suspended while we were waiting?
if (!jt->handle_special_suspend_equivalent_condition()) break ;
// This thread was externally suspended
//
// This logic isn't needed for JVMTI raw monitors,
// but doesn't hurt just in case the suspend rules change. This
// logic is needed for the ObjectMonitor.wait() reentry phase.
// We have reentered the contended monitor, but while we were
// waiting another thread suspended us. We don't want to reenter
// the monitor while suspended because that would surprise the
// thread that suspended us.
//
// Drop the lock -
SimpleExit (THREAD) ;
jt->java_suspend_self();
}
assert(_owner == THREAD, "Fatal error with monitor owner!");
assert(_recursions == 0, "Fatal error with monitor recursions!");
}
THREAD->set_current_pending_monitor(NULL);
guarantee (_recursions == 0, "invariant") ;
return OM_OK;
}
// Used mainly for JVMTI raw monitor implementation
// Also used for ObjectMonitor::wait().
int ObjectMonitor::raw_exit(TRAPS) {
TEVENT (raw_exit) ;
if (THREAD != _owner) {
return OM_ILLEGAL_MONITOR_STATE;
}
if (_recursions > 0) {
--_recursions ;
return OM_OK ;
}
void * List = _EntryList ;
SimpleExit (THREAD) ;
return OM_OK;
}
// Used for JVMTI raw monitor implementation.
// All JavaThreads will enter here with state _thread_blocked
int ObjectMonitor::raw_wait(jlong millis, bool interruptible, TRAPS) {
TEVENT (raw_wait) ;
if (THREAD != _owner) {
return OM_ILLEGAL_MONITOR_STATE;
}
// To avoid spurious wakeups we reset the parkevent -- This is strictly optional.
// The caller must be able to tolerate spurious returns from raw_wait().
THREAD->_ParkEvent->reset() ;
OrderAccess::fence() ;
// check interrupt event
if (interruptible && Thread::is_interrupted(THREAD, true)) {
return OM_INTERRUPTED;
}
intptr_t save = _recursions ;
_recursions = 0 ;
_waiters ++ ;
if (THREAD->is_Java_thread()) {
guarantee (((JavaThread *) THREAD)->thread_state() == _thread_blocked, "invariant") ;
((JavaThread *)THREAD)->set_suspend_equivalent();
}
int rv = SimpleWait (THREAD, millis) ;
_recursions = save ;
_waiters -- ;
guarantee (THREAD == _owner, "invariant") ;
if (THREAD->is_Java_thread()) {
JavaThread * jSelf = (JavaThread *) THREAD ;
for (;;) {
if (!jSelf->handle_special_suspend_equivalent_condition()) break ;
SimpleExit (THREAD) ;
jSelf->java_suspend_self();
SimpleEnter (THREAD) ;
jSelf->set_suspend_equivalent() ;
}
}
guarantee (THREAD == _owner, "invariant") ;
if (interruptible && Thread::is_interrupted(THREAD, true)) {
return OM_INTERRUPTED;
}
return OM_OK ;
}
int ObjectMonitor::raw_notify(TRAPS) {
TEVENT (raw_notify) ;
if (THREAD != _owner) {
return OM_ILLEGAL_MONITOR_STATE;
}
SimpleNotify (THREAD, false) ;
return OM_OK;
}
int ObjectMonitor::raw_notifyAll(TRAPS) {
TEVENT (raw_notifyAll) ;
if (THREAD != _owner) {
return OM_ILLEGAL_MONITOR_STATE;
}
SimpleNotify (THREAD, true) ;
return OM_OK;
}
#ifndef PRODUCT
void ObjectMonitor::verify() {
}
void ObjectMonitor::print() {
}
#endif
//------------------------------------------------------------------------------
// Non-product code
#ifndef PRODUCT
void ObjectSynchronizer::trace_locking(Handle locking_obj, bool is_compiled,
bool is_method, bool is_locking) {
// Don't know what to do here
}
// Verify all monitors in the monitor cache, the verification is weak.
void ObjectSynchronizer::verify() {
ObjectMonitor* block = gBlockList;
ObjectMonitor* mid;
while (block) {
assert(block->object() == CHAINMARKER, "must be a block header");
for (int i = 1; i < _BLOCKSIZE; i++) {
mid = block + i;
oop object = (oop) mid->object();
if (object != NULL) {
mid->verify();
}
}
block = (ObjectMonitor*) block->FreeNext;
}
}
// Check if monitor belongs to the monitor cache
// The list is grow-only so it's *relatively* safe to traverse
// the list of extant blocks without taking a lock.
int ObjectSynchronizer::verify_objmon_isinpool(ObjectMonitor *monitor) {
ObjectMonitor* block = gBlockList;
while (block) {
assert(block->object() == CHAINMARKER, "must be a block header");
if (monitor > &block[0] && monitor < &block[_BLOCKSIZE]) {
address mon = (address) monitor;
address blk = (address) block;
size_t diff = mon - blk;
assert((diff % sizeof(ObjectMonitor)) == 0, "check");
return 1;
}
block = (ObjectMonitor*) block->FreeNext;
}
return 0;
}
#endif
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