| 105 |
|
* successful atomic operation per enq/deq pair. But it also |
| 106 |
|
* enables lower cost variants of queue maintenance mechanics. (A |
| 107 |
|
* variation of this idea applies even for non-dual queues that |
| 108 |
< |
* support deletion of embedded elements, such as |
| 108 |
> |
* support deletion of interior elements, such as |
| 109 |
|
* j.u.c.ConcurrentLinkedQueue.) |
| 110 |
|
* |
| 111 |
< |
* Once a node is matched, its item can never again change. We |
| 112 |
< |
* may thus arrange that the linked list of them contains a prefix |
| 113 |
< |
* of zero or more matched nodes, followed by a suffix of zero or |
| 114 |
< |
* more unmatched nodes. (Note that we allow both the prefix and |
| 115 |
< |
* suffix to be zero length, which in turn means that we do not |
| 116 |
< |
* use a dummy header.) If we were not concerned with either time |
| 117 |
< |
* or space efficiency, we could correctly perform enqueue and |
| 118 |
< |
* dequeue operations by traversing from a pointer to the initial |
| 119 |
< |
* node; CASing the item of the first unmatched node on match and |
| 120 |
< |
* CASing the next field of the trailing node on appends. While |
| 121 |
< |
* this would be a terrible idea in itself, it does have the |
| 122 |
< |
* benefit of not requiring ANY atomic updates on head/tail |
| 123 |
< |
* fields. |
| 111 |
> |
* Once a node is matched, its match status can never again |
| 112 |
> |
* change. We may thus arrange that the linked list of them |
| 113 |
> |
* contain a prefix of zero or more matched nodes, followed by a |
| 114 |
> |
* suffix of zero or more unmatched nodes. (Note that we allow |
| 115 |
> |
* both the prefix and suffix to be zero length, which in turn |
| 116 |
> |
* means that we do not use a dummy header.) If we were not |
| 117 |
> |
* concerned with either time or space efficiency, we could |
| 118 |
> |
* correctly perform enqueue and dequeue operations by traversing |
| 119 |
> |
* from a pointer to the initial node; CASing the item of the |
| 120 |
> |
* first unmatched node on match and CASing the next field of the |
| 121 |
> |
* trailing node on appends. (Plus some special-casing when |
| 122 |
> |
* initially empty). While this would be a terrible idea in |
| 123 |
> |
* itself, it does have the benefit of not requiring ANY atomic |
| 124 |
> |
* updates on head/tail fields. |
| 125 |
|
* |
| 126 |
|
* We introduce here an approach that lies between the extremes of |
| 127 |
< |
* never versus always updating queue (head and tail) pointers |
| 128 |
< |
* that reflects the tradeoff of sometimes requiring extra traversal |
| 129 |
< |
* steps to locate the first and/or last unmatched nodes, versus |
| 130 |
< |
* the reduced overhead and contention of fewer updates to queue |
| 131 |
< |
* pointers. For example, a possible snapshot of a queue is: |
| 127 |
> |
* never versus always updating queue (head and tail) pointers. |
| 128 |
> |
* This offers a tradeoff between sometimes requiring extra |
| 129 |
> |
* traversal steps to locate the first and/or last unmatched |
| 130 |
> |
* nodes, versus the reduced overhead and contention of fewer |
| 131 |
> |
* updates to queue pointers. For example, a possible snapshot of |
| 132 |
> |
* a queue is: |
| 133 |
|
* |
| 134 |
|
* head tail |
| 135 |
|
* | | |
| 141 |
|
* similarly for "tail") is an empirical matter. We have found |
| 142 |
|
* that using very small constants in the range of 1-3 work best |
| 143 |
|
* over a range of platforms. Larger values introduce increasing |
| 144 |
< |
* costs of cache misses and risks of long traversal chains. |
| 144 |
> |
* costs of cache misses and risks of long traversal chains, while |
| 145 |
> |
* smaller values increase CAS contention and overhead. |
| 146 |
|
* |
| 147 |
|
* Dual queues with slack differ from plain M&S dual queues by |
| 148 |
|
* virtue of only sometimes updating head or tail pointers when |
| 161 |
|
* targets. Even when using very small slack values, this |
| 162 |
|
* approach works well for dual queues because it allows all |
| 163 |
|
* operations up to the point of matching or appending an item |
| 164 |
< |
* (hence potentially releasing another thread) to be read-only, |
| 165 |
< |
* thus not introducing any further contention. As described |
| 166 |
< |
* below, we implement this by performing slack maintenance |
| 167 |
< |
* retries only after these points. |
| 164 |
> |
* (hence potentially allowing progress by another thread) to be |
| 165 |
> |
* read-only, thus not introducing any further contention. As |
| 166 |
> |
* described below, we implement this by performing slack |
| 167 |
> |
* maintenance retries only after these points. |
| 168 |
|
* |
| 169 |
|
* As an accompaniment to such techniques, traversal overhead can |
| 170 |
|
* be further reduced without increasing contention of head |
| 171 |
< |
* pointer updates. During traversals, threads may sometimes |
| 172 |
< |
* shortcut the "next" link path from the current "head" node to |
| 173 |
< |
* be closer to the currently known first unmatched node. Again, |
| 174 |
< |
* this may be triggered with using thresholds or randomization. |
| 171 |
> |
* pointer updates: Threads may sometimes shortcut the "next" link |
| 172 |
> |
* path from the current "head" node to be closer to the currently |
| 173 |
> |
* known first unmatched node, and similarly for tail. Again, this |
| 174 |
> |
* may be triggered with using thresholds or randomization. |
| 175 |
|
* |
| 176 |
|
* These ideas must be further extended to avoid unbounded amounts |
| 177 |
|
* of costly-to-reclaim garbage caused by the sequential "next" |
| 199 |
|
* mechanics because an update may leave head at a detached node. |
| 200 |
|
* And while direct writes are possible for tail updates, they |
| 201 |
|
* increase the risk of long retraversals, and hence long garbage |
| 202 |
< |
* chains which can be much more costly than is worthwhile |
| 202 |
> |
* chains, which can be much more costly than is worthwhile |
| 203 |
|
* considering that the cost difference of performing a CAS vs |
| 204 |
|
* write is smaller when they are not triggered on each operation |
| 205 |
|
* (especially considering that writes and CASes equally require |
| 206 |
|
* additional GC bookkeeping ("write barriers") that are sometimes |
| 207 |
|
* more costly than the writes themselves because of contention). |
| 208 |
|
* |
| 209 |
< |
* Removal of internal nodes (due to timed out or interrupted |
| 210 |
< |
* waits, or calls to remove or Iterator.remove) uses a scheme |
| 211 |
< |
* roughly similar to that in Scherer, Lea, and Scott |
| 212 |
< |
* SynchronousQueue. Given a predecessor, we can unsplice any node |
| 213 |
< |
* except the (actual) tail of the queue. To avoid build-up of |
| 214 |
< |
* cancelled trailing nodes, upon a request to remove a trailing |
| 215 |
< |
* node, it is placed in field "cleanMe" to be unspliced later. |
| 209 |
> |
* Removal of interior nodes (due to timed out or interrupted |
| 210 |
> |
* waits, or calls to remove(x) or Iterator.remove) can use a |
| 211 |
> |
* scheme roughly similar to that described in Scherer, Lea, and |
| 212 |
> |
* Scott's SynchronousQueue. Given a predecessor, we can unsplice |
| 213 |
> |
* any node except the (actual) tail of the queue. To avoid |
| 214 |
> |
* build-up of cancelled trailing nodes, upon a request to remove |
| 215 |
> |
* a trailing node, it is placed in field "cleanMe" to be |
| 216 |
> |
* unspliced upon the next call to unsplice any other node. |
| 217 |
> |
* Situations needing such mechanics are not common but do occur |
| 218 |
> |
* in practice; for example when an unbounded series of short |
| 219 |
> |
* timed calls to poll repeatedly time out but never otherwise |
| 220 |
> |
* fall off the list because of an untimed call to take at the |
| 221 |
> |
* front of the queue. Note that maintaining field cleanMe does |
| 222 |
> |
* not otherwise much impact garbage retention even if never |
| 223 |
> |
* cleared by some other call because the held node will |
| 224 |
> |
* eventually either directly or indirectly lead to a self-link |
| 225 |
> |
* once off the list. |
| 226 |
|
* |
| 227 |
|
* *** Overview of implementation *** |
| 228 |
|
* |
| 229 |
< |
* We use a threshold-based approach to updates, with a target |
| 230 |
< |
* slack of two. The slack value is hard-wired: a path greater |
| 229 |
> |
* We use a threshold-based approach to updates, with a slack |
| 230 |
> |
* threshold of two -- that is, we update head/tail when the |
| 231 |
> |
* current pointer appears to be two or more steps away from the |
| 232 |
> |
* first/last node. The slack value is hard-wired: a path greater |
| 233 |
|
* than one is naturally implemented by checking equality of |
| 234 |
|
* traversal pointers except when the list has only one element, |
| 235 |
< |
* in which case we keep max slack at one. Avoiding tracking |
| 236 |
< |
* explicit counts across situations slightly simplifies an |
| 235 |
> |
* in which case we keep slack threshold at one. Avoiding tracking |
| 236 |
> |
* explicit counts across method calls slightly simplifies an |
| 237 |
|
* already-messy implementation. Using randomization would |
| 238 |
|
* probably work better if there were a low-quality dirt-cheap |
| 239 |
|
* per-thread one available, but even ThreadLocalRandom is too |
| 240 |
|
* heavy for these purposes. |
| 241 |
|
* |
| 242 |
< |
* With such a small slack value, path short-circuiting is rarely |
| 243 |
< |
* worthwhile. However, it is used (in awaitMatch) immediately |
| 244 |
< |
* before a waiting thread starts to block, as a final bit of |
| 245 |
< |
* helping at a point when contention with others is extremely |
| 246 |
< |
* unlikely (since if other threads that could release it are |
| 247 |
< |
* operating, then the current thread wouldn't be blocking). |
| 242 |
> |
* With such a small slack threshold value, it is rarely |
| 243 |
> |
* worthwhile to augment this with path short-circuiting; i.e., |
| 244 |
> |
* unsplicing nodes between head and the first unmatched node, or |
| 245 |
> |
* similarly for tail, rather than advancing head or tail |
| 246 |
> |
* proper. However, it is used (in awaitMatch) immediately before |
| 247 |
> |
* a waiting thread starts to block, as a final bit of helping at |
| 248 |
> |
* a point when contention with others is extremely unlikely |
| 249 |
> |
* (since if other threads that could release it are operating, |
| 250 |
> |
* then the current thread wouldn't be blocking). |
| 251 |
> |
* |
| 252 |
> |
* We allow both the head and tail fields to be null before any |
| 253 |
> |
* nodes are enqueued; initializing upon first append. This |
| 254 |
> |
* simplifies some other logic, as well as providing more |
| 255 |
> |
* efficient explicit control paths instead of letting JVMs insert |
| 256 |
> |
* implicit NullPointerExceptions when they are null. While not |
| 257 |
> |
* currently fully implemented, we also leave open the possibility |
| 258 |
> |
* of re-nulling these fields when empty (which is complicated to |
| 259 |
> |
* arrange, for little benefit.) |
| 260 |
|
* |
| 261 |
|
* All enqueue/dequeue operations are handled by the single method |
| 262 |
|
* "xfer" with parameters indicating whether to act as some form |
| 263 |
|
* of offer, put, poll, take, or transfer (each possibly with |
| 264 |
|
* timeout). The relative complexity of using one monolithic |
| 265 |
|
* method outweighs the code bulk and maintenance problems of |
| 266 |
< |
* using nine separate methods. |
| 266 |
> |
* using separate methods for each case. |
| 267 |
|
* |
| 268 |
|
* Operation consists of up to three phases. The first is |
| 269 |
|
* implemented within method xfer, the second in tryAppend, and |
| 276 |
|
* case matching it and returning, also if necessary updating |
| 277 |
|
* head to one past the matched node (or the node itself if the |
| 278 |
|
* list has no other unmatched nodes). If the CAS misses, then |
| 279 |
< |
* a retry loops until the slack is at most two. Traversals |
| 280 |
< |
* also check if the initial head is now off-list, in which |
| 281 |
< |
* case they start at the new head. |
| 279 |
> |
* a loop retries advancing head by two steps until either |
| 280 |
> |
* success or the slack is at most two. By requiring that each |
| 281 |
> |
* attempt advances head by two (if applicable), we ensure that |
| 282 |
> |
* the slack does not grow without bound. Traversals also check |
| 283 |
> |
* if the initial head is now off-list, in which case they |
| 284 |
> |
* start at the new head. |
| 285 |
|
* |
| 286 |
|
* If no candidates are found and the call was untimed |
| 287 |
|
* poll/offer, (argument "how" is NOW) return. |
| 288 |
|
* |
| 289 |
|
* 2. Try to append a new node (method tryAppend) |
| 290 |
|
* |
| 291 |
< |
* Starting at current tail pointer, try to append a new node |
| 292 |
< |
* to the list (or if head was null, establish the first |
| 293 |
< |
* node). Nodes can be appended only if their predecessors are |
| 294 |
< |
* either already matched or are of the same mode. If we detect |
| 295 |
< |
* otherwise, then a new node with opposite mode must have been |
| 296 |
< |
* appended during traversal, so must restart at phase 1. The |
| 297 |
< |
* traversal and update steps are otherwise similar to phase 1: |
| 298 |
< |
* Retrying upon CAS misses and checking for staleness. In |
| 299 |
< |
* particular, if a self-link is encountered, then we can |
| 300 |
< |
* safely jump to a node on the list by continuing the |
| 301 |
< |
* traversal at current head. |
| 291 |
> |
* Starting at current tail pointer, find the actual last node |
| 292 |
> |
* and try to append a new node (or if head was null, establish |
| 293 |
> |
* the first node). Nodes can be appended only if their |
| 294 |
> |
* predecessors are either already matched or are of the same |
| 295 |
> |
* mode. If we detect otherwise, then a new node with opposite |
| 296 |
> |
* mode must have been appended during traversal, so we must |
| 297 |
> |
* restart at phase 1. The traversal and update steps are |
| 298 |
> |
* otherwise similar to phase 1: Retrying upon CAS misses and |
| 299 |
> |
* checking for staleness. In particular, if a self-link is |
| 300 |
> |
* encountered, then we can safely jump to a node on the list |
| 301 |
> |
* by continuing the traversal at current head. |
| 302 |
|
* |
| 303 |
|
* On successful append, if the call was ASYNC, return. |
| 304 |
|
* |
| 305 |
|
* 3. Await match or cancellation (method awaitMatch) |
| 306 |
|
* |
| 307 |
|
* Wait for another thread to match node; instead cancelling if |
| 308 |
< |
* current thread was interrupted or the wait timed out. On |
| 308 |
> |
* the current thread was interrupted or the wait timed out. On |
| 309 |
|
* multiprocessors, we use front-of-queue spinning: If a node |
| 310 |
|
* appears to be the first unmatched node in the queue, it |
| 311 |
|
* spins a bit before blocking. In either case, before blocking |
| 320 |
|
* to decide to occasionally perform a Thread.yield. While |
| 321 |
|
* yield has underdefined specs, we assume that might it help, |
| 322 |
|
* and will not hurt in limiting impact of spinning on busy |
| 323 |
< |
* systems. We also use much smaller (1/4) spins for nodes |
| 324 |
< |
* that are not known to be front but whose predecessors have |
| 325 |
< |
* not blocked -- these "chained" spins avoid artifacts of |
| 323 |
> |
* systems. We also use smaller (1/2) spins for nodes that are |
| 324 |
> |
* not known to be front but whose predecessors have not |
| 325 |
> |
* blocked -- these "chained" spins avoid artifacts of |
| 326 |
|
* front-of-queue rules which otherwise lead to alternating |
| 327 |
|
* nodes spinning vs blocking. Further, front threads that |
| 328 |
|
* represent phase changes (from data to request node or vice |
| 329 |
|
* versa) compared to their predecessors receive additional |
| 330 |
< |
* spins, reflecting the longer code path lengths necessary to |
| 331 |
< |
* release them under contention. |
| 330 |
> |
* chained spins, reflecting longer paths typically required to |
| 331 |
> |
* unblock threads during phase changes. |
| 332 |
|
*/ |
| 333 |
|
|
| 334 |
|
/** True if on multiprocessor */ |
| 336 |
|
Runtime.getRuntime().availableProcessors() > 1; |
| 337 |
|
|
| 338 |
|
/** |
| 339 |
< |
* The number of times to spin (with on average one randomly |
| 340 |
< |
* interspersed call to Thread.yield) on multiprocessor before |
| 341 |
< |
* blocking when a node is apparently the first waiter in the |
| 342 |
< |
* queue. See above for explanation. Must be a power of two. The |
| 343 |
< |
* value is empirically derived -- it works pretty well across a |
| 344 |
< |
* variety of processors, numbers of CPUs, and OSes. |
| 339 |
> |
* The number of times to spin (with randomly interspersed calls |
| 340 |
> |
* to Thread.yield) on multiprocessor before blocking when a node |
| 341 |
> |
* is apparently the first waiter in the queue. See above for |
| 342 |
> |
* explanation. Must be a power of two. The value is empirically |
| 343 |
> |
* derived -- it works pretty well across a variety of processors, |
| 344 |
> |
* numbers of CPUs, and OSes. |
| 345 |
|
*/ |
| 346 |
|
private static final int FRONT_SPINS = 1 << 7; |
| 347 |
|
|
| 348 |
|
/** |
| 349 |
|
* The number of times to spin before blocking when a node is |
| 350 |
< |
* preceded by another node that is apparently spinning. |
| 350 |
> |
* preceded by another node that is apparently spinning. Also |
| 351 |
> |
* serves as an increment to FRONT_SPINS on phase changes, and as |
| 352 |
> |
* base average frequency for yielding during spins. Must be a |
| 353 |
> |
* power of two. |
| 354 |
|
*/ |
| 355 |
< |
private static final int CHAINED_SPINS = FRONT_SPINS >>> 2; |
| 355 |
> |
private static final int CHAINED_SPINS = FRONT_SPINS >>> 1; |
| 356 |
|
|
| 357 |
|
/** |
| 358 |
|
* Queue nodes. Uses Object, not E, for items to allow forgetting |
| 502 |
|
if (isData == haveData) // can't match |
| 503 |
|
break; |
| 504 |
|
if (p.casItem(item, e)) { // match |
| 505 |
< |
Thread w = p.waiter; |
| 506 |
< |
while (p != h) { // update head |
| 507 |
< |
Node n = p.next; // by 2 unless singleton |
| 508 |
< |
if (n != null) |
| 509 |
< |
p = n; |
| 477 |
< |
if (head == h && casHead(h, p)) { |
| 505 |
> |
for (Node q = p; q != h;) { |
| 506 |
> |
Node n = q.next; // update head by 2 |
| 507 |
> |
if (n != null) // unless singleton |
| 508 |
> |
q = n; |
| 509 |
> |
if (head == h && casHead(h, q)) { |
| 510 |
|
h.forgetNext(); |
| 511 |
|
break; |
| 512 |
|
} // advance and retry |
| 513 |
|
if ((h = head) == null || |
| 514 |
< |
(p = h.next) == null || !p.isMatched()) |
| 514 |
> |
(q = h.next) == null || !q.isMatched()) |
| 515 |
|
break; // unless slack < 2 |
| 516 |
|
} |
| 517 |
< |
LockSupport.unpark(w); |
| 517 |
> |
LockSupport.unpark(p.waiter); |
| 518 |
|
return item; |
| 519 |
|
} |
| 520 |
|
} |
| 529 |
|
if (pred == null) |
| 530 |
|
continue retry; // lost race vs opposite mode |
| 531 |
|
if (how >= SYNC) |
| 532 |
< |
return awaitMatch(pred, s, e, how, nanos); |
| 532 |
> |
return awaitMatch(s, pred, e, how, nanos); |
| 533 |
|
} |
| 534 |
|
return e; // not waiting |
| 535 |
|
} |
| 538 |
|
/** |
| 539 |
|
* Tries to append node s as tail. |
| 540 |
|
* |
| 509 |
– |
* @param haveData true if appending in data mode |
| 541 |
|
* @param s the node to append |
| 542 |
+ |
* @param haveData true if appending in data mode |
| 543 |
|
* @return null on failure due to losing race with append in |
| 544 |
|
* different mode, else s's predecessor, or s itself if no |
| 545 |
|
* predecessor |
| 546 |
|
*/ |
| 547 |
|
private Node tryAppend(Node s, boolean haveData) { |
| 548 |
< |
for (Node t = tail, p = t;;) { // move p to actual tail and append |
| 548 |
> |
for (Node t = tail, p = t;;) { // move p to last node and append |
| 549 |
|
Node n, u; // temps for reads of next & tail |
| 550 |
|
if (p == null && (p = head) == null) { |
| 551 |
|
if (casHead(null, s)) |
| 553 |
|
} |
| 554 |
|
else if (p.cannotPrecede(haveData)) |
| 555 |
|
return null; // lost race vs opposite mode |
| 556 |
< |
else if ((n = p.next) != null) // Not tail; keep traversing |
| 556 |
> |
else if ((n = p.next) != null) // not last; keep traversing |
| 557 |
|
p = p != t && t != (u = tail) ? (t = u) : // stale tail |
| 558 |
|
(p != n) ? n : null; // restart if off list |
| 559 |
|
else if (!p.casNext(null, s)) |
| 560 |
|
p = p.next; // re-read on CAS failure |
| 561 |
|
else { |
| 562 |
< |
if (p != t) { // Update if slack now >= 2 |
| 562 |
> |
if (p != t) { // update if slack now >= 2 |
| 563 |
|
while ((tail != t || !casTail(t, s)) && |
| 564 |
|
(t = tail) != null && |
| 565 |
|
(s = t.next) != null && // advance and retry |
| 573 |
|
/** |
| 574 |
|
* Spins/yields/blocks until node s is matched or caller gives up. |
| 575 |
|
* |
| 544 |
– |
* @param pred the predecessor of s or s or null if none |
| 576 |
|
* @param s the waiting node |
| 577 |
+ |
* @param pred the predecessor of s, or s itself if it has no |
| 578 |
+ |
* predecessor, or null if unknown (the null case does not occur |
| 579 |
+ |
* in any current calls but may in possible future extensions) |
| 580 |
|
* @param e the comparison value for checking match |
| 581 |
|
* @param how either SYNC or TIMEOUT |
| 582 |
|
* @param nanos timeout value |
| 583 |
|
* @return matched item, or e if unmatched on interrupt or timeout |
| 584 |
|
*/ |
| 585 |
< |
private Object awaitMatch(Node pred, Node s, Object e, |
| 585 |
> |
private Object awaitMatch(Node s, Node pred, Object e, |
| 586 |
|
int how, long nanos) { |
| 587 |
|
long lastTime = (how == TIMEOUT) ? System.nanoTime() : 0L; |
| 588 |
|
Thread w = Thread.currentThread(); |
| 605 |
|
if ((spins = spinsFor(pred, s.isData)) > 0) |
| 606 |
|
randomYields = ThreadLocalRandom.current(); |
| 607 |
|
} |
| 608 |
< |
else if (spins > 0) { // spin, occasionally yield |
| 609 |
< |
if (randomYields.nextInt(FRONT_SPINS) == 0) |
| 610 |
< |
Thread.yield(); |
| 611 |
< |
--spins; |
| 608 |
> |
else if (spins > 0) { // spin |
| 609 |
> |
if (--spins == 0) |
| 610 |
> |
shortenHeadPath(); // reduce slack before blocking |
| 611 |
> |
else if (randomYields.nextInt(CHAINED_SPINS) == 0) |
| 612 |
> |
Thread.yield(); // occasionally yield |
| 613 |
|
} |
| 614 |
|
else if (s.waiter == null) { |
| 615 |
< |
shortenHeadPath(); // reduce slack before blocking |
| 581 |
< |
s.waiter = w; // request unpark |
| 615 |
> |
s.waiter = w; // request unpark then recheck |
| 616 |
|
} |
| 617 |
|
else if (how == TIMEOUT) { |
| 618 |
|
long now = System.nanoTime(); |
| 622 |
|
} |
| 623 |
|
else { |
| 624 |
|
LockSupport.park(this); |
| 625 |
+ |
s.waiter = null; |
| 626 |
|
spins = -1; // spin if front upon wakeup |
| 627 |
|
} |
| 628 |
|
} |
| 634 |
|
*/ |
| 635 |
|
private static int spinsFor(Node pred, boolean haveData) { |
| 636 |
|
if (MP && pred != null) { |
| 637 |
< |
boolean predData = pred.isData; |
| 638 |
< |
if (predData != haveData) // front and phase change |
| 639 |
< |
return FRONT_SPINS + (FRONT_SPINS >>> 1); |
| 605 |
< |
if (predData != (pred.item != null)) // probably at front |
| 637 |
> |
if (pred.isData != haveData) // phase change |
| 638 |
> |
return FRONT_SPINS + CHAINED_SPINS; |
| 639 |
> |
if (pred.isMatched()) // probably at front |
| 640 |
|
return FRONT_SPINS; |
| 641 |
|
if (pred.waiter == null) // pred apparently spinning |
| 642 |
|
return CHAINED_SPINS; |
| 788 |
|
s.forgetContents(); // clear unneeded fields |
| 789 |
|
/* |
| 790 |
|
* At any given time, exactly one node on list cannot be |
| 791 |
< |
* deleted -- the last inserted node. To accommodate this, if |
| 792 |
< |
* we cannot delete s, we save its predecessor as "cleanMe", |
| 791 |
> |
* unlinked -- the last inserted node. To accommodate this, if |
| 792 |
> |
* we cannot unlink s, we save its predecessor as "cleanMe", |
| 793 |
|
* processing the previously saved version first. Because only |
| 794 |
|
* one node in the list can have a null next, at least one of |
| 795 |
|
* node s or the node previously saved can always be |
| 1192 |
|
} |
| 1193 |
|
} |
| 1194 |
|
|
| 1161 |
– |
|
| 1195 |
|
// Unsafe mechanics |
| 1196 |
|
|
| 1197 |
|
private static final sun.misc.Unsafe UNSAFE = getUnsafe(); |
| 1214 |
|
} |
| 1215 |
|
} |
| 1216 |
|
|
| 1217 |
+ |
/** |
| 1218 |
+ |
* Returns a sun.misc.Unsafe. Suitable for use in a 3rd party package. |
| 1219 |
+ |
* Replace with a simple call to Unsafe.getUnsafe when integrating |
| 1220 |
+ |
* into a jdk. |
| 1221 |
+ |
* |
| 1222 |
+ |
* @return a sun.misc.Unsafe |
| 1223 |
+ |
*/ |
| 1224 |
|
private static sun.misc.Unsafe getUnsafe() { |
| 1225 |
|
try { |
| 1226 |
|
return sun.misc.Unsafe.getUnsafe(); |
