| /* |
| * Written by Doug Lea with assistance from members of JCP JSR-166 |
| * Expert Group and released to the public domain, as explained at |
| * http://creativecommons.org/publicdomain/zero/1.0/ |
| */ |
| |
| package java.util.concurrent; |
| |
| import java.util.AbstractQueue; |
| import java.util.Collection; |
| import java.util.Iterator; |
| import java.util.NoSuchElementException; |
| import java.util.Queue; |
| import java.util.concurrent.TimeUnit; |
| import java.util.concurrent.locks.LockSupport; |
| |
| // BEGIN android-note |
| // removed link to collections framework docs |
| // END android-note |
| |
| /** |
| * An unbounded {@link TransferQueue} based on linked nodes. |
| * This queue orders elements FIFO (first-in-first-out) with respect |
| * to any given producer. The <em>head</em> of the queue is that |
| * element that has been on the queue the longest time for some |
| * producer. The <em>tail</em> of the queue is that element that has |
| * been on the queue the shortest time for some producer. |
| * |
| * <p>Beware that, unlike in most collections, the {@code size} method |
| * is <em>NOT</em> a constant-time operation. Because of the |
| * asynchronous nature of these queues, determining the current number |
| * of elements requires a traversal of the elements, and so may report |
| * inaccurate results if this collection is modified during traversal. |
| * Additionally, the bulk operations {@code addAll}, |
| * {@code removeAll}, {@code retainAll}, {@code containsAll}, |
| * {@code equals}, and {@code toArray} are <em>not</em> guaranteed |
| * to be performed atomically. For example, an iterator operating |
| * concurrently with an {@code addAll} operation might view only some |
| * of the added elements. |
| * |
| * <p>This class and its iterator implement all of the |
| * <em>optional</em> methods of the {@link Collection} and {@link |
| * Iterator} interfaces. |
| * |
| * <p>Memory consistency effects: As with other concurrent |
| * collections, actions in a thread prior to placing an object into a |
| * {@code LinkedTransferQueue} |
| * <a href="package-summary.html#MemoryVisibility"><i>happen-before</i></a> |
| * actions subsequent to the access or removal of that element from |
| * the {@code LinkedTransferQueue} in another thread. |
| * |
| * @since 1.7 |
| * @hide |
| * @author Doug Lea |
| * @param <E> the type of elements held in this collection |
| */ |
| public class LinkedTransferQueue<E> extends AbstractQueue<E> |
| implements TransferQueue<E>, java.io.Serializable { |
| private static final long serialVersionUID = -3223113410248163686L; |
| |
| /* |
| * *** Overview of Dual Queues with Slack *** |
| * |
| * Dual Queues, introduced by Scherer and Scott |
| * (http://www.cs.rice.edu/~wns1/papers/2004-DISC-DDS.pdf) are |
| * (linked) queues in which nodes may represent either data or |
| * requests. When a thread tries to enqueue a data node, but |
| * encounters a request node, it instead "matches" and removes it; |
| * and vice versa for enqueuing requests. Blocking Dual Queues |
| * arrange that threads enqueuing unmatched requests block until |
| * other threads provide the match. Dual Synchronous Queues (see |
| * Scherer, Lea, & Scott |
| * http://www.cs.rochester.edu/u/scott/papers/2009_Scherer_CACM_SSQ.pdf) |
| * additionally arrange that threads enqueuing unmatched data also |
| * block. Dual Transfer Queues support all of these modes, as |
| * dictated by callers. |
| * |
| * A FIFO dual queue may be implemented using a variation of the |
| * Michael & Scott (M&S) lock-free queue algorithm |
| * (http://www.cs.rochester.edu/u/scott/papers/1996_PODC_queues.pdf). |
| * It maintains two pointer fields, "head", pointing to a |
| * (matched) node that in turn points to the first actual |
| * (unmatched) queue node (or null if empty); and "tail" that |
| * points to the last node on the queue (or again null if |
| * empty). For example, here is a possible queue with four data |
| * elements: |
| * |
| * head tail |
| * | | |
| * v v |
| * M -> U -> U -> U -> U |
| * |
| * The M&S queue algorithm is known to be prone to scalability and |
| * overhead limitations when maintaining (via CAS) these head and |
| * tail pointers. This has led to the development of |
| * contention-reducing variants such as elimination arrays (see |
| * Moir et al http://portal.acm.org/citation.cfm?id=1074013) and |
| * optimistic back pointers (see Ladan-Mozes & Shavit |
| * http://people.csail.mit.edu/edya/publications/OptimisticFIFOQueue-journal.pdf). |
| * However, the nature of dual queues enables a simpler tactic for |
| * improving M&S-style implementations when dual-ness is needed. |
| * |
| * In a dual queue, each node must atomically maintain its match |
| * status. While there are other possible variants, we implement |
| * this here as: for a data-mode node, matching entails CASing an |
| * "item" field from a non-null data value to null upon match, and |
| * vice-versa for request nodes, CASing from null to a data |
| * value. (Note that the linearization properties of this style of |
| * queue are easy to verify -- elements are made available by |
| * linking, and unavailable by matching.) Compared to plain M&S |
| * queues, this property of dual queues requires one additional |
| * successful atomic operation per enq/deq pair. But it also |
| * enables lower cost variants of queue maintenance mechanics. (A |
| * variation of this idea applies even for non-dual queues that |
| * support deletion of interior elements, such as |
| * j.u.c.ConcurrentLinkedQueue.) |
| * |
| * Once a node is matched, its match status can never again |
| * change. We may thus arrange that the linked list of them |
| * contain a prefix of zero or more matched nodes, followed by a |
| * suffix of zero or more unmatched nodes. (Note that we allow |
| * both the prefix and suffix to be zero length, which in turn |
| * means that we do not use a dummy header.) If we were not |
| * concerned with either time or space efficiency, we could |
| * correctly perform enqueue and dequeue operations by traversing |
| * from a pointer to the initial node; CASing the item of the |
| * first unmatched node on match and CASing the next field of the |
| * trailing node on appends. (Plus some special-casing when |
| * initially empty). While this would be a terrible idea in |
| * itself, it does have the benefit of not requiring ANY atomic |
| * updates on head/tail fields. |
| * |
| * We introduce here an approach that lies between the extremes of |
| * never versus always updating queue (head and tail) pointers. |
| * This offers a tradeoff between sometimes requiring extra |
| * traversal steps to locate the first and/or last unmatched |
| * nodes, versus the reduced overhead and contention of fewer |
| * updates to queue pointers. For example, a possible snapshot of |
| * a queue is: |
| * |
| * head tail |
| * | | |
| * v v |
| * M -> M -> U -> U -> U -> U |
| * |
| * The best value for this "slack" (the targeted maximum distance |
| * between the value of "head" and the first unmatched node, and |
| * similarly for "tail") is an empirical matter. We have found |
| * that using very small constants in the range of 1-3 work best |
| * over a range of platforms. Larger values introduce increasing |
| * costs of cache misses and risks of long traversal chains, while |
| * smaller values increase CAS contention and overhead. |
| * |
| * Dual queues with slack differ from plain M&S dual queues by |
| * virtue of only sometimes updating head or tail pointers when |
| * matching, appending, or even traversing nodes; in order to |
| * maintain a targeted slack. The idea of "sometimes" may be |
| * operationalized in several ways. The simplest is to use a |
| * per-operation counter incremented on each traversal step, and |
| * to try (via CAS) to update the associated queue pointer |
| * whenever the count exceeds a threshold. Another, that requires |
| * more overhead, is to use random number generators to update |
| * with a given probability per traversal step. |
| * |
| * In any strategy along these lines, because CASes updating |
| * fields may fail, the actual slack may exceed targeted |
| * slack. However, they may be retried at any time to maintain |
| * targets. Even when using very small slack values, this |
| * approach works well for dual queues because it allows all |
| * operations up to the point of matching or appending an item |
| * (hence potentially allowing progress by another thread) to be |
| * read-only, thus not introducing any further contention. As |
| * described below, we implement this by performing slack |
| * maintenance retries only after these points. |
| * |
| * As an accompaniment to such techniques, traversal overhead can |
| * be further reduced without increasing contention of head |
| * pointer updates: Threads may sometimes shortcut the "next" link |
| * path from the current "head" node to be closer to the currently |
| * known first unmatched node, and similarly for tail. Again, this |
| * may be triggered with using thresholds or randomization. |
| * |
| * These ideas must be further extended to avoid unbounded amounts |
| * of costly-to-reclaim garbage caused by the sequential "next" |
| * links of nodes starting at old forgotten head nodes: As first |
| * described in detail by Boehm |
| * (http://portal.acm.org/citation.cfm?doid=503272.503282) if a GC |
| * delays noticing that any arbitrarily old node has become |
| * garbage, all newer dead nodes will also be unreclaimed. |
| * (Similar issues arise in non-GC environments.) To cope with |
| * this in our implementation, upon CASing to advance the head |
| * pointer, we set the "next" link of the previous head to point |
| * only to itself; thus limiting the length of connected dead lists. |
| * (We also take similar care to wipe out possibly garbage |
| * retaining values held in other Node fields.) However, doing so |
| * adds some further complexity to traversal: If any "next" |
| * pointer links to itself, it indicates that the current thread |
| * has lagged behind a head-update, and so the traversal must |
| * continue from the "head". Traversals trying to find the |
| * current tail starting from "tail" may also encounter |
| * self-links, in which case they also continue at "head". |
| * |
| * It is tempting in slack-based scheme to not even use CAS for |
| * updates (similarly to Ladan-Mozes & Shavit). However, this |
| * cannot be done for head updates under the above link-forgetting |
| * mechanics because an update may leave head at a detached node. |
| * And while direct writes are possible for tail updates, they |
| * increase the risk of long retraversals, and hence long garbage |
| * chains, which can be much more costly than is worthwhile |
| * considering that the cost difference of performing a CAS vs |
| * write is smaller when they are not triggered on each operation |
| * (especially considering that writes and CASes equally require |
| * additional GC bookkeeping ("write barriers") that are sometimes |
| * more costly than the writes themselves because of contention). |
| * |
| * *** Overview of implementation *** |
| * |
| * We use a threshold-based approach to updates, with a slack |
| * threshold of two -- that is, we update head/tail when the |
| * current pointer appears to be two or more steps away from the |
| * first/last node. The slack value is hard-wired: a path greater |
| * than one is naturally implemented by checking equality of |
| * traversal pointers except when the list has only one element, |
| * in which case we keep slack threshold at one. Avoiding tracking |
| * explicit counts across method calls slightly simplifies an |
| * already-messy implementation. Using randomization would |
| * probably work better if there were a low-quality dirt-cheap |
| * per-thread one available, but even ThreadLocalRandom is too |
| * heavy for these purposes. |
| * |
| * With such a small slack threshold value, it is not worthwhile |
| * to augment this with path short-circuiting (i.e., unsplicing |
| * interior nodes) except in the case of cancellation/removal (see |
| * below). |
| * |
| * We allow both the head and tail fields to be null before any |
| * nodes are enqueued; initializing upon first append. This |
| * simplifies some other logic, as well as providing more |
| * efficient explicit control paths instead of letting JVMs insert |
| * implicit NullPointerExceptions when they are null. While not |
| * currently fully implemented, we also leave open the possibility |
| * of re-nulling these fields when empty (which is complicated to |
| * arrange, for little benefit.) |
| * |
| * All enqueue/dequeue operations are handled by the single method |
| * "xfer" with parameters indicating whether to act as some form |
| * of offer, put, poll, take, or transfer (each possibly with |
| * timeout). The relative complexity of using one monolithic |
| * method outweighs the code bulk and maintenance problems of |
| * using separate methods for each case. |
| * |
| * Operation consists of up to three phases. The first is |
| * implemented within method xfer, the second in tryAppend, and |
| * the third in method awaitMatch. |
| * |
| * 1. Try to match an existing node |
| * |
| * Starting at head, skip already-matched nodes until finding |
| * an unmatched node of opposite mode, if one exists, in which |
| * case matching it and returning, also if necessary updating |
| * head to one past the matched node (or the node itself if the |
| * list has no other unmatched nodes). If the CAS misses, then |
| * a loop retries advancing head by two steps until either |
| * success or the slack is at most two. By requiring that each |
| * attempt advances head by two (if applicable), we ensure that |
| * the slack does not grow without bound. Traversals also check |
| * if the initial head is now off-list, in which case they |
| * start at the new head. |
| * |
| * If no candidates are found and the call was untimed |
| * poll/offer, (argument "how" is NOW) return. |
| * |
| * 2. Try to append a new node (method tryAppend) |
| * |
| * Starting at current tail pointer, find the actual last node |
| * and try to append a new node (or if head was null, establish |
| * the first node). Nodes can be appended only if their |
| * predecessors are either already matched or are of the same |
| * mode. If we detect otherwise, then a new node with opposite |
| * mode must have been appended during traversal, so we must |
| * restart at phase 1. The traversal and update steps are |
| * otherwise similar to phase 1: Retrying upon CAS misses and |
| * checking for staleness. In particular, if a self-link is |
| * encountered, then we can safely jump to a node on the list |
| * by continuing the traversal at current head. |
| * |
| * On successful append, if the call was ASYNC, return. |
| * |
| * 3. Await match or cancellation (method awaitMatch) |
| * |
| * Wait for another thread to match node; instead cancelling if |
| * the current thread was interrupted or the wait timed out. On |
| * multiprocessors, we use front-of-queue spinning: If a node |
| * appears to be the first unmatched node in the queue, it |
| * spins a bit before blocking. In either case, before blocking |
| * it tries to unsplice any nodes between the current "head" |
| * and the first unmatched node. |
| * |
| * Front-of-queue spinning vastly improves performance of |
| * heavily contended queues. And so long as it is relatively |
| * brief and "quiet", spinning does not much impact performance |
| * of less-contended queues. During spins threads check their |
| * interrupt status and generate a thread-local random number |
| * to decide to occasionally perform a Thread.yield. While |
| * yield has underdefined specs, we assume that it might help, |
| * and will not hurt, in limiting impact of spinning on busy |
| * systems. We also use smaller (1/2) spins for nodes that are |
| * not known to be front but whose predecessors have not |
| * blocked -- these "chained" spins avoid artifacts of |
| * front-of-queue rules which otherwise lead to alternating |
| * nodes spinning vs blocking. Further, front threads that |
| * represent phase changes (from data to request node or vice |
| * versa) compared to their predecessors receive additional |
| * chained spins, reflecting longer paths typically required to |
| * unblock threads during phase changes. |
| * |
| * |
| * ** Unlinking removed interior nodes ** |
| * |
| * In addition to minimizing garbage retention via self-linking |
| * described above, we also unlink removed interior nodes. These |
| * may arise due to timed out or interrupted waits, or calls to |
| * remove(x) or Iterator.remove. Normally, given a node that was |
| * at one time known to be the predecessor of some node s that is |
| * to be removed, we can unsplice s by CASing the next field of |
| * its predecessor if it still points to s (otherwise s must |
| * already have been removed or is now offlist). But there are two |
| * situations in which we cannot guarantee to make node s |
| * unreachable in this way: (1) If s is the trailing node of list |
| * (i.e., with null next), then it is pinned as the target node |
| * for appends, so can only be removed later after other nodes are |
| * appended. (2) We cannot necessarily unlink s given a |
| * predecessor node that is matched (including the case of being |
| * cancelled): the predecessor may already be unspliced, in which |
| * case some previous reachable node may still point to s. |
| * (For further explanation see Herlihy & Shavit "The Art of |
| * Multiprocessor Programming" chapter 9). Although, in both |
| * cases, we can rule out the need for further action if either s |
| * or its predecessor are (or can be made to be) at, or fall off |
| * from, the head of list. |
| * |
| * Without taking these into account, it would be possible for an |
| * unbounded number of supposedly removed nodes to remain |
| * reachable. Situations leading to such buildup are uncommon but |
| * can occur in practice; for example when a series of short timed |
| * calls to poll repeatedly time out but never otherwise fall off |
| * the list because of an untimed call to take at the front of the |
| * queue. |
| * |
| * When these cases arise, rather than always retraversing the |
| * entire list to find an actual predecessor to unlink (which |
| * won't help for case (1) anyway), we record a conservative |
| * estimate of possible unsplice failures (in "sweepVotes"). |
| * We trigger a full sweep when the estimate exceeds a threshold |
| * ("SWEEP_THRESHOLD") indicating the maximum number of estimated |
| * removal failures to tolerate before sweeping through, unlinking |
| * cancelled nodes that were not unlinked upon initial removal. |
| * We perform sweeps by the thread hitting threshold (rather than |
| * background threads or by spreading work to other threads) |
| * because in the main contexts in which removal occurs, the |
| * caller is already timed-out, cancelled, or performing a |
| * potentially O(n) operation (e.g. remove(x)), none of which are |
| * time-critical enough to warrant the overhead that alternatives |
| * would impose on other threads. |
| * |
| * Because the sweepVotes estimate is conservative, and because |
| * nodes become unlinked "naturally" as they fall off the head of |
| * the queue, and because we allow votes to accumulate even while |
| * sweeps are in progress, there are typically significantly fewer |
| * such nodes than estimated. Choice of a threshold value |
| * balances the likelihood of wasted effort and contention, versus |
| * providing a worst-case bound on retention of interior nodes in |
| * quiescent queues. The value defined below was chosen |
| * empirically to balance these under various timeout scenarios. |
| * |
| * Note that we cannot self-link unlinked interior nodes during |
| * sweeps. However, the associated garbage chains terminate when |
| * some successor ultimately falls off the head of the list and is |
| * self-linked. |
| */ |
| |
| /** True if on multiprocessor */ |
| private static final boolean MP = |
| Runtime.getRuntime().availableProcessors() > 1; |
| |
| /** |
| * The number of times to spin (with randomly interspersed calls |
| * to Thread.yield) on multiprocessor before blocking when a node |
| * is apparently the first waiter in the queue. See above for |
| * explanation. Must be a power of two. The value is empirically |
| * derived -- it works pretty well across a variety of processors, |
| * numbers of CPUs, and OSes. |
| */ |
| private static final int FRONT_SPINS = 1 << 7; |
| |
| /** |
| * The number of times to spin before blocking when a node is |
| * preceded by another node that is apparently spinning. Also |
| * serves as an increment to FRONT_SPINS on phase changes, and as |
| * base average frequency for yielding during spins. Must be a |
| * power of two. |
| */ |
| private static final int CHAINED_SPINS = FRONT_SPINS >>> 1; |
| |
| /** |
| * The maximum number of estimated removal failures (sweepVotes) |
| * to tolerate before sweeping through the queue unlinking |
| * cancelled nodes that were not unlinked upon initial |
| * removal. See above for explanation. The value must be at least |
| * two to avoid useless sweeps when removing trailing nodes. |
| */ |
| static final int SWEEP_THRESHOLD = 32; |
| |
| /** |
| * Queue nodes. Uses Object, not E, for items to allow forgetting |
| * them after use. Relies heavily on Unsafe mechanics to minimize |
| * unnecessary ordering constraints: Writes that are intrinsically |
| * ordered wrt other accesses or CASes use simple relaxed forms. |
| */ |
| static final class Node { |
| final boolean isData; // false if this is a request node |
| volatile Object item; // initially non-null if isData; CASed to match |
| volatile Node next; |
| volatile Thread waiter; // null until waiting |
| |
| // CAS methods for fields |
| final boolean casNext(Node cmp, Node val) { |
| return UNSAFE.compareAndSwapObject(this, nextOffset, cmp, val); |
| } |
| |
| final boolean casItem(Object cmp, Object val) { |
| // assert cmp == null || cmp.getClass() != Node.class; |
| return UNSAFE.compareAndSwapObject(this, itemOffset, cmp, val); |
| } |
| |
| /** |
| * Constructs a new node. Uses relaxed write because item can |
| * only be seen after publication via casNext. |
| */ |
| Node(Object item, boolean isData) { |
| UNSAFE.putObject(this, itemOffset, item); // relaxed write |
| this.isData = isData; |
| } |
| |
| /** |
| * Links node to itself to avoid garbage retention. Called |
| * only after CASing head field, so uses relaxed write. |
| */ |
| final void forgetNext() { |
| UNSAFE.putObject(this, nextOffset, this); |
| } |
| |
| /** |
| * Sets item to self and waiter to null, to avoid garbage |
| * retention after matching or cancelling. Uses relaxed writes |
| * because order is already constrained in the only calling |
| * contexts: item is forgotten only after volatile/atomic |
| * mechanics that extract items. Similarly, clearing waiter |
| * follows either CAS or return from park (if ever parked; |
| * else we don't care). |
| */ |
| final void forgetContents() { |
| UNSAFE.putObject(this, itemOffset, this); |
| UNSAFE.putObject(this, waiterOffset, null); |
| } |
| |
| /** |
| * Returns true if this node has been matched, including the |
| * case of artificial matches due to cancellation. |
| */ |
| final boolean isMatched() { |
| Object x = item; |
| return (x == this) || ((x == null) == isData); |
| } |
| |
| /** |
| * Returns true if this is an unmatched request node. |
| */ |
| final boolean isUnmatchedRequest() { |
| return !isData && item == null; |
| } |
| |
| /** |
| * Returns true if a node with the given mode cannot be |
| * appended to this node because this node is unmatched and |
| * has opposite data mode. |
| */ |
| final boolean cannotPrecede(boolean haveData) { |
| boolean d = isData; |
| Object x; |
| return d != haveData && (x = item) != this && (x != null) == d; |
| } |
| |
| /** |
| * Tries to artificially match a data node -- used by remove. |
| */ |
| final boolean tryMatchData() { |
| // assert isData; |
| Object x = item; |
| if (x != null && x != this && casItem(x, null)) { |
| LockSupport.unpark(waiter); |
| return true; |
| } |
| return false; |
| } |
| |
| private static final long serialVersionUID = -3375979862319811754L; |
| |
| // Unsafe mechanics |
| private static final sun.misc.Unsafe UNSAFE; |
| private static final long itemOffset; |
| private static final long nextOffset; |
| private static final long waiterOffset; |
| static { |
| try { |
| UNSAFE = sun.misc.Unsafe.getUnsafe(); |
| Class<?> k = Node.class; |
| itemOffset = UNSAFE.objectFieldOffset |
| (k.getDeclaredField("item")); |
| nextOffset = UNSAFE.objectFieldOffset |
| (k.getDeclaredField("next")); |
| waiterOffset = UNSAFE.objectFieldOffset |
| (k.getDeclaredField("waiter")); |
| } catch (Exception e) { |
| throw new Error(e); |
| } |
| } |
| } |
| |
| /** head of the queue; null until first enqueue */ |
| transient volatile Node head; |
| |
| /** tail of the queue; null until first append */ |
| private transient volatile Node tail; |
| |
| /** The number of apparent failures to unsplice removed nodes */ |
| private transient volatile int sweepVotes; |
| |
| // CAS methods for fields |
| private boolean casTail(Node cmp, Node val) { |
| return UNSAFE.compareAndSwapObject(this, tailOffset, cmp, val); |
| } |
| |
| private boolean casHead(Node cmp, Node val) { |
| return UNSAFE.compareAndSwapObject(this, headOffset, cmp, val); |
| } |
| |
| private boolean casSweepVotes(int cmp, int val) { |
| return UNSAFE.compareAndSwapInt(this, sweepVotesOffset, cmp, val); |
| } |
| |
| /* |
| * Possible values for "how" argument in xfer method. |
| */ |
| private static final int NOW = 0; // for untimed poll, tryTransfer |
| private static final int ASYNC = 1; // for offer, put, add |
| private static final int SYNC = 2; // for transfer, take |
| private static final int TIMED = 3; // for timed poll, tryTransfer |
| |
| @SuppressWarnings("unchecked") |
| static <E> E cast(Object item) { |
| // assert item == null || item.getClass() != Node.class; |
| return (E) item; |
| } |
| |
| /** |
| * Implements all queuing methods. See above for explanation. |
| * |
| * @param e the item or null for take |
| * @param haveData true if this is a put, else a take |
| * @param how NOW, ASYNC, SYNC, or TIMED |
| * @param nanos timeout in nanosecs, used only if mode is TIMED |
| * @return an item if matched, else e |
| * @throws NullPointerException if haveData mode but e is null |
| */ |
| private E xfer(E e, boolean haveData, int how, long nanos) { |
| if (haveData && (e == null)) |
| throw new NullPointerException(); |
| Node s = null; // the node to append, if needed |
| |
| retry: |
| for (;;) { // restart on append race |
| |
| for (Node h = head, p = h; p != null;) { // find & match first node |
| boolean isData = p.isData; |
| Object item = p.item; |
| if (item != p && (item != null) == isData) { // unmatched |
| if (isData == haveData) // can't match |
| break; |
| if (p.casItem(item, e)) { // match |
| for (Node q = p; q != h;) { |
| Node n = q.next; // update by 2 unless singleton |
| if (head == h && casHead(h, n == null ? q : n)) { |
| h.forgetNext(); |
| break; |
| } // advance and retry |
| if ((h = head) == null || |
| (q = h.next) == null || !q.isMatched()) |
| break; // unless slack < 2 |
| } |
| LockSupport.unpark(p.waiter); |
| return LinkedTransferQueue.<E>cast(item); |
| } |
| } |
| Node n = p.next; |
| p = (p != n) ? n : (h = head); // Use head if p offlist |
| } |
| |
| if (how != NOW) { // No matches available |
| if (s == null) |
| s = new Node(e, haveData); |
| Node pred = tryAppend(s, haveData); |
| if (pred == null) |
| continue retry; // lost race vs opposite mode |
| if (how != ASYNC) |
| return awaitMatch(s, pred, e, (how == TIMED), nanos); |
| } |
| return e; // not waiting |
| } |
| } |
| |
| /** |
| * Tries to append node s as tail. |
| * |
| * @param s the node to append |
| * @param haveData true if appending in data mode |
| * @return null on failure due to losing race with append in |
| * different mode, else s's predecessor, or s itself if no |
| * predecessor |
| */ |
| private Node tryAppend(Node s, boolean haveData) { |
| for (Node t = tail, p = t;;) { // move p to last node and append |
| Node n, u; // temps for reads of next & tail |
| if (p == null && (p = head) == null) { |
| if (casHead(null, s)) |
| return s; // initialize |
| } |
| else if (p.cannotPrecede(haveData)) |
| return null; // lost race vs opposite mode |
| else if ((n = p.next) != null) // not last; keep traversing |
| p = p != t && t != (u = tail) ? (t = u) : // stale tail |
| (p != n) ? n : null; // restart if off list |
| else if (!p.casNext(null, s)) |
| p = p.next; // re-read on CAS failure |
| else { |
| if (p != t) { // update if slack now >= 2 |
| while ((tail != t || !casTail(t, s)) && |
| (t = tail) != null && |
| (s = t.next) != null && // advance and retry |
| (s = s.next) != null && s != t); |
| } |
| return p; |
| } |
| } |
| } |
| |
| /** |
| * Spins/yields/blocks until node s is matched or caller gives up. |
| * |
| * @param s the waiting node |
| * @param pred the predecessor of s, or s itself if it has no |
| * predecessor, or null if unknown (the null case does not occur |
| * in any current calls but may in possible future extensions) |
| * @param e the comparison value for checking match |
| * @param timed if true, wait only until timeout elapses |
| * @param nanos timeout in nanosecs, used only if timed is true |
| * @return matched item, or e if unmatched on interrupt or timeout |
| */ |
| private E awaitMatch(Node s, Node pred, E e, boolean timed, long nanos) { |
| long lastTime = timed ? System.nanoTime() : 0L; |
| Thread w = Thread.currentThread(); |
| int spins = -1; // initialized after first item and cancel checks |
| ThreadLocalRandom randomYields = null; // bound if needed |
| |
| for (;;) { |
| Object item = s.item; |
| if (item != e) { // matched |
| // assert item != s; |
| s.forgetContents(); // avoid garbage |
| return LinkedTransferQueue.<E>cast(item); |
| } |
| if ((w.isInterrupted() || (timed && nanos <= 0)) && |
| s.casItem(e, s)) { // cancel |
| unsplice(pred, s); |
| return e; |
| } |
| |
| if (spins < 0) { // establish spins at/near front |
| if ((spins = spinsFor(pred, s.isData)) > 0) |
| randomYields = ThreadLocalRandom.current(); |
| } |
| else if (spins > 0) { // spin |
| --spins; |
| if (randomYields.nextInt(CHAINED_SPINS) == 0) |
| Thread.yield(); // occasionally yield |
| } |
| else if (s.waiter == null) { |
| s.waiter = w; // request unpark then recheck |
| } |
| else if (timed) { |
| long now = System.nanoTime(); |
| if ((nanos -= now - lastTime) > 0) |
| LockSupport.parkNanos(this, nanos); |
| lastTime = now; |
| } |
| else { |
| LockSupport.park(this); |
| } |
| } |
| } |
| |
| /** |
| * Returns spin/yield value for a node with given predecessor and |
| * data mode. See above for explanation. |
| */ |
| private static int spinsFor(Node pred, boolean haveData) { |
| if (MP && pred != null) { |
| if (pred.isData != haveData) // phase change |
| return FRONT_SPINS + CHAINED_SPINS; |
| if (pred.isMatched()) // probably at front |
| return FRONT_SPINS; |
| if (pred.waiter == null) // pred apparently spinning |
| return CHAINED_SPINS; |
| } |
| return 0; |
| } |
| |
| /* -------------- Traversal methods -------------- */ |
| |
| /** |
| * Returns the successor of p, or the head node if p.next has been |
| * linked to self, which will only be true if traversing with a |
| * stale pointer that is now off the list. |
| */ |
| final Node succ(Node p) { |
| Node next = p.next; |
| return (p == next) ? head : next; |
| } |
| |
| /** |
| * Returns the first unmatched node of the given mode, or null if |
| * none. Used by methods isEmpty, hasWaitingConsumer. |
| */ |
| private Node firstOfMode(boolean isData) { |
| for (Node p = head; p != null; p = succ(p)) { |
| if (!p.isMatched()) |
| return (p.isData == isData) ? p : null; |
| } |
| return null; |
| } |
| |
| /** |
| * Returns the item in the first unmatched node with isData; or |
| * null if none. Used by peek. |
| */ |
| private E firstDataItem() { |
| for (Node p = head; p != null; p = succ(p)) { |
| Object item = p.item; |
| if (p.isData) { |
| if (item != null && item != p) |
| return LinkedTransferQueue.<E>cast(item); |
| } |
| else if (item == null) |
| return null; |
| } |
| return null; |
| } |
| |
| /** |
| * Traverses and counts unmatched nodes of the given mode. |
| * Used by methods size and getWaitingConsumerCount. |
| */ |
| private int countOfMode(boolean data) { |
| int count = 0; |
| for (Node p = head; p != null; ) { |
| if (!p.isMatched()) { |
| if (p.isData != data) |
| return 0; |
| if (++count == Integer.MAX_VALUE) // saturated |
| break; |
| } |
| Node n = p.next; |
| if (n != p) |
| p = n; |
| else { |
| count = 0; |
| p = head; |
| } |
| } |
| return count; |
| } |
| |
| final class Itr implements Iterator<E> { |
| private Node nextNode; // next node to return item for |
| private E nextItem; // the corresponding item |
| private Node lastRet; // last returned node, to support remove |
| private Node lastPred; // predecessor to unlink lastRet |
| |
| /** |
| * Moves to next node after prev, or first node if prev null. |
| */ |
| private void advance(Node prev) { |
| /* |
| * To track and avoid buildup of deleted nodes in the face |
| * of calls to both Queue.remove and Itr.remove, we must |
| * include variants of unsplice and sweep upon each |
| * advance: Upon Itr.remove, we may need to catch up links |
| * from lastPred, and upon other removes, we might need to |
| * skip ahead from stale nodes and unsplice deleted ones |
| * found while advancing. |
| */ |
| |
| Node r, b; // reset lastPred upon possible deletion of lastRet |
| if ((r = lastRet) != null && !r.isMatched()) |
| lastPred = r; // next lastPred is old lastRet |
| else if ((b = lastPred) == null || b.isMatched()) |
| lastPred = null; // at start of list |
| else { |
| Node s, n; // help with removal of lastPred.next |
| while ((s = b.next) != null && |
| s != b && s.isMatched() && |
| (n = s.next) != null && n != s) |
| b.casNext(s, n); |
| } |
| |
| this.lastRet = prev; |
| |
| for (Node p = prev, s, n;;) { |
| s = (p == null) ? head : p.next; |
| if (s == null) |
| break; |
| else if (s == p) { |
| p = null; |
| continue; |
| } |
| Object item = s.item; |
| if (s.isData) { |
| if (item != null && item != s) { |
| nextItem = LinkedTransferQueue.<E>cast(item); |
| nextNode = s; |
| return; |
| } |
| } |
| else if (item == null) |
| break; |
| // assert s.isMatched(); |
| if (p == null) |
| p = s; |
| else if ((n = s.next) == null) |
| break; |
| else if (s == n) |
| p = null; |
| else |
| p.casNext(s, n); |
| } |
| nextNode = null; |
| nextItem = null; |
| } |
| |
| Itr() { |
| advance(null); |
| } |
| |
| public final boolean hasNext() { |
| return nextNode != null; |
| } |
| |
| public final E next() { |
| Node p = nextNode; |
| if (p == null) throw new NoSuchElementException(); |
| E e = nextItem; |
| advance(p); |
| return e; |
| } |
| |
| public final void remove() { |
| final Node lastRet = this.lastRet; |
| if (lastRet == null) |
| throw new IllegalStateException(); |
| this.lastRet = null; |
| if (lastRet.tryMatchData()) |
| unsplice(lastPred, lastRet); |
| } |
| } |
| |
| /* -------------- Removal methods -------------- */ |
| |
| /** |
| * Unsplices (now or later) the given deleted/cancelled node with |
| * the given predecessor. |
| * |
| * @param pred a node that was at one time known to be the |
| * predecessor of s, or null or s itself if s is/was at head |
| * @param s the node to be unspliced |
| */ |
| final void unsplice(Node pred, Node s) { |
| s.forgetContents(); // forget unneeded fields |
| /* |
| * See above for rationale. Briefly: if pred still points to |
| * s, try to unlink s. If s cannot be unlinked, because it is |
| * trailing node or pred might be unlinked, and neither pred |
| * nor s are head or offlist, add to sweepVotes, and if enough |
| * votes have accumulated, sweep. |
| */ |
| if (pred != null && pred != s && pred.next == s) { |
| Node n = s.next; |
| if (n == null || |
| (n != s && pred.casNext(s, n) && pred.isMatched())) { |
| for (;;) { // check if at, or could be, head |
| Node h = head; |
| if (h == pred || h == s || h == null) |
| return; // at head or list empty |
| if (!h.isMatched()) |
| break; |
| Node hn = h.next; |
| if (hn == null) |
| return; // now empty |
| if (hn != h && casHead(h, hn)) |
| h.forgetNext(); // advance head |
| } |
| if (pred.next != pred && s.next != s) { // recheck if offlist |
| for (;;) { // sweep now if enough votes |
| int v = sweepVotes; |
| if (v < SWEEP_THRESHOLD) { |
| if (casSweepVotes(v, v + 1)) |
| break; |
| } |
| else if (casSweepVotes(v, 0)) { |
| sweep(); |
| break; |
| } |
| } |
| } |
| } |
| } |
| } |
| |
| /** |
| * Unlinks matched (typically cancelled) nodes encountered in a |
| * traversal from head. |
| */ |
| private void sweep() { |
| for (Node p = head, s, n; p != null && (s = p.next) != null; ) { |
| if (!s.isMatched()) |
| // Unmatched nodes are never self-linked |
| p = s; |
| else if ((n = s.next) == null) // trailing node is pinned |
| break; |
| else if (s == n) // stale |
| // No need to also check for p == s, since that implies s == n |
| p = head; |
| else |
| p.casNext(s, n); |
| } |
| } |
| |
| /** |
| * Main implementation of remove(Object) |
| */ |
| private boolean findAndRemove(Object e) { |
| if (e != null) { |
| for (Node pred = null, p = head; p != null; ) { |
| Object item = p.item; |
| if (p.isData) { |
| if (item != null && item != p && e.equals(item) && |
| p.tryMatchData()) { |
| unsplice(pred, p); |
| return true; |
| } |
| } |
| else if (item == null) |
| break; |
| pred = p; |
| if ((p = p.next) == pred) { // stale |
| pred = null; |
| p = head; |
| } |
| } |
| } |
| return false; |
| } |
| |
| |
| /** |
| * Creates an initially empty {@code LinkedTransferQueue}. |
| */ |
| public LinkedTransferQueue() { |
| } |
| |
| /** |
| * Creates a {@code LinkedTransferQueue} |
| * initially containing the elements of the given collection, |
| * added in traversal order of the collection's iterator. |
| * |
| * @param c the collection of elements to initially contain |
| * @throws NullPointerException if the specified collection or any |
| * of its elements are null |
| */ |
| public LinkedTransferQueue(Collection<? extends E> c) { |
| this(); |
| addAll(c); |
| } |
| |
| /** |
| * Inserts the specified element at the tail of this queue. |
| * As the queue is unbounded, this method will never block. |
| * |
| * @throws NullPointerException if the specified element is null |
| */ |
| public void put(E e) { |
| xfer(e, true, ASYNC, 0); |
| } |
| |
| /** |
| * Inserts the specified element at the tail of this queue. |
| * As the queue is unbounded, this method will never block or |
| * return {@code false}. |
| * |
| * @return {@code true} (as specified by |
| * {@link java.util.concurrent.BlockingQueue#offer(Object,long,TimeUnit) |
| * BlockingQueue.offer}) |
| * @throws NullPointerException if the specified element is null |
| */ |
| public boolean offer(E e, long timeout, TimeUnit unit) { |
| xfer(e, true, ASYNC, 0); |
| return true; |
| } |
| |
| /** |
| * Inserts the specified element at the tail of this queue. |
| * As the queue is unbounded, this method will never return {@code false}. |
| * |
| * @return {@code true} (as specified by {@link Queue#offer}) |
| * @throws NullPointerException if the specified element is null |
| */ |
| public boolean offer(E e) { |
| xfer(e, true, ASYNC, 0); |
| return true; |
| } |
| |
| /** |
| * Inserts the specified element at the tail of this queue. |
| * As the queue is unbounded, this method will never throw |
| * {@link IllegalStateException} or return {@code false}. |
| * |
| * @return {@code true} (as specified by {@link Collection#add}) |
| * @throws NullPointerException if the specified element is null |
| */ |
| public boolean add(E e) { |
| xfer(e, true, ASYNC, 0); |
| return true; |
| } |
| |
| /** |
| * Transfers the element to a waiting consumer immediately, if possible. |
| * |
| * <p>More precisely, transfers the specified element immediately |
| * if there exists a consumer already waiting to receive it (in |
| * {@link #take} or timed {@link #poll(long,TimeUnit) poll}), |
| * otherwise returning {@code false} without enqueuing the element. |
| * |
| * @throws NullPointerException if the specified element is null |
| */ |
| public boolean tryTransfer(E e) { |
| return xfer(e, true, NOW, 0) == null; |
| } |
| |
| /** |
| * Transfers the element to a consumer, waiting if necessary to do so. |
| * |
| * <p>More precisely, transfers the specified element immediately |
| * if there exists a consumer already waiting to receive it (in |
| * {@link #take} or timed {@link #poll(long,TimeUnit) poll}), |
| * else inserts the specified element at the tail of this queue |
| * and waits until the element is received by a consumer. |
| * |
| * @throws NullPointerException if the specified element is null |
| */ |
| public void transfer(E e) throws InterruptedException { |
| if (xfer(e, true, SYNC, 0) != null) { |
| Thread.interrupted(); // failure possible only due to interrupt |
| throw new InterruptedException(); |
| } |
| } |
| |
| /** |
| * Transfers the element to a consumer if it is possible to do so |
| * before the timeout elapses. |
| * |
| * <p>More precisely, transfers the specified element immediately |
| * if there exists a consumer already waiting to receive it (in |
| * {@link #take} or timed {@link #poll(long,TimeUnit) poll}), |
| * else inserts the specified element at the tail of this queue |
| * and waits until the element is received by a consumer, |
| * returning {@code false} if the specified wait time elapses |
| * before the element can be transferred. |
| * |
| * @throws NullPointerException if the specified element is null |
| */ |
| public boolean tryTransfer(E e, long timeout, TimeUnit unit) |
| throws InterruptedException { |
| if (xfer(e, true, TIMED, unit.toNanos(timeout)) == null) |
| return true; |
| if (!Thread.interrupted()) |
| return false; |
| throw new InterruptedException(); |
| } |
| |
| public E take() throws InterruptedException { |
| E e = xfer(null, false, SYNC, 0); |
| if (e != null) |
| return e; |
| Thread.interrupted(); |
| throw new InterruptedException(); |
| } |
| |
| public E poll(long timeout, TimeUnit unit) throws InterruptedException { |
| E e = xfer(null, false, TIMED, unit.toNanos(timeout)); |
| if (e != null || !Thread.interrupted()) |
| return e; |
| throw new InterruptedException(); |
| } |
| |
| public E poll() { |
| return xfer(null, false, NOW, 0); |
| } |
| |
| /** |
| * @throws NullPointerException {@inheritDoc} |
| * @throws IllegalArgumentException {@inheritDoc} |
| */ |
| public int drainTo(Collection<? super E> c) { |
| if (c == null) |
| throw new NullPointerException(); |
| if (c == this) |
| throw new IllegalArgumentException(); |
| int n = 0; |
| for (E e; (e = poll()) != null;) { |
| c.add(e); |
| ++n; |
| } |
| return n; |
| } |
| |
| /** |
| * @throws NullPointerException {@inheritDoc} |
| * @throws IllegalArgumentException {@inheritDoc} |
| */ |
| public int drainTo(Collection<? super E> c, int maxElements) { |
| if (c == null) |
| throw new NullPointerException(); |
| if (c == this) |
| throw new IllegalArgumentException(); |
| int n = 0; |
| for (E e; n < maxElements && (e = poll()) != null;) { |
| c.add(e); |
| ++n; |
| } |
| return n; |
| } |
| |
| /** |
| * Returns an iterator over the elements in this queue in proper sequence. |
| * The elements will be returned in order from first (head) to last (tail). |
| * |
| * <p>The returned iterator is a "weakly consistent" iterator that |
| * will never throw {@link java.util.ConcurrentModificationException |
| * ConcurrentModificationException}, and guarantees to traverse |
| * elements as they existed upon construction of the iterator, and |
| * may (but is not guaranteed to) reflect any modifications |
| * subsequent to construction. |
| * |
| * @return an iterator over the elements in this queue in proper sequence |
| */ |
| public Iterator<E> iterator() { |
| return new Itr(); |
| } |
| |
| public E peek() { |
| return firstDataItem(); |
| } |
| |
| /** |
| * Returns {@code true} if this queue contains no elements. |
| * |
| * @return {@code true} if this queue contains no elements |
| */ |
| public boolean isEmpty() { |
| for (Node p = head; p != null; p = succ(p)) { |
| if (!p.isMatched()) |
| return !p.isData; |
| } |
| return true; |
| } |
| |
| public boolean hasWaitingConsumer() { |
| return firstOfMode(false) != null; |
| } |
| |
| /** |
| * Returns the number of elements in this queue. If this queue |
| * contains more than {@code Integer.MAX_VALUE} elements, returns |
| * {@code Integer.MAX_VALUE}. |
| * |
| * <p>Beware that, unlike in most collections, this method is |
| * <em>NOT</em> a constant-time operation. Because of the |
| * asynchronous nature of these queues, determining the current |
| * number of elements requires an O(n) traversal. |
| * |
| * @return the number of elements in this queue |
| */ |
| public int size() { |
| return countOfMode(true); |
| } |
| |
| public int getWaitingConsumerCount() { |
| return countOfMode(false); |
| } |
| |
| /** |
| * Removes a single instance of the specified element from this queue, |
| * if it is present. More formally, removes an element {@code e} such |
| * that {@code o.equals(e)}, if this queue contains one or more such |
| * elements. |
| * Returns {@code true} if this queue contained the specified element |
| * (or equivalently, if this queue changed as a result of the call). |
| * |
| * @param o element to be removed from this queue, if present |
| * @return {@code true} if this queue changed as a result of the call |
| */ |
| public boolean remove(Object o) { |
| return findAndRemove(o); |
| } |
| |
| /** |
| * Returns {@code true} if this queue contains the specified element. |
| * More formally, returns {@code true} if and only if this queue contains |
| * at least one element {@code e} such that {@code o.equals(e)}. |
| * |
| * @param o object to be checked for containment in this queue |
| * @return {@code true} if this queue contains the specified element |
| */ |
| public boolean contains(Object o) { |
| if (o == null) return false; |
| for (Node p = head; p != null; p = succ(p)) { |
| Object item = p.item; |
| if (p.isData) { |
| if (item != null && item != p && o.equals(item)) |
| return true; |
| } |
| else if (item == null) |
| break; |
| } |
| return false; |
| } |
| |
| /** |
| * Always returns {@code Integer.MAX_VALUE} because a |
| * {@code LinkedTransferQueue} is not capacity constrained. |
| * |
| * @return {@code Integer.MAX_VALUE} (as specified by |
| * {@link java.util.concurrent.BlockingQueue#remainingCapacity() |
| * BlockingQueue.remainingCapacity}) |
| */ |
| public int remainingCapacity() { |
| return Integer.MAX_VALUE; |
| } |
| |
| /** |
| * Saves the state to a stream (that is, serializes it). |
| * |
| * @serialData All of the elements (each an {@code E}) in |
| * the proper order, followed by a null |
| * @param s the stream |
| */ |
| private void writeObject(java.io.ObjectOutputStream s) |
| throws java.io.IOException { |
| s.defaultWriteObject(); |
| for (E e : this) |
| s.writeObject(e); |
| // Use trailing null as sentinel |
| s.writeObject(null); |
| } |
| |
| /** |
| * Reconstitutes the Queue instance from a stream (that is, |
| * deserializes it). |
| * |
| * @param s the stream |
| */ |
| private void readObject(java.io.ObjectInputStream s) |
| throws java.io.IOException, ClassNotFoundException { |
| s.defaultReadObject(); |
| for (;;) { |
| @SuppressWarnings("unchecked") E item = (E) s.readObject(); |
| if (item == null) |
| break; |
| else |
| offer(item); |
| } |
| } |
| |
| // Unsafe mechanics |
| |
| private static final sun.misc.Unsafe UNSAFE; |
| private static final long headOffset; |
| private static final long tailOffset; |
| private static final long sweepVotesOffset; |
| static { |
| try { |
| UNSAFE = sun.misc.Unsafe.getUnsafe(); |
| Class<?> k = LinkedTransferQueue.class; |
| headOffset = UNSAFE.objectFieldOffset |
| (k.getDeclaredField("head")); |
| tailOffset = UNSAFE.objectFieldOffset |
| (k.getDeclaredField("tail")); |
| sweepVotesOffset = UNSAFE.objectFieldOffset |
| (k.getDeclaredField("sweepVotes")); |
| } catch (Exception e) { |
| throw new Error(e); |
| } |
| } |
| } |
