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| author | Akinori Ito <aito@eie.yz.yamagata-u.ac.jp> | 2001-11-15 00:32:13 +0000 |
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| committer | Akinori Ito <aito@eie.yz.yamagata-u.ac.jp> | 2001-11-15 00:32:13 +0000 |
| commit | 85da7ee692072c643939e9f4b24fbd1e74e64e70 (patch) | |
| tree | 9fc63298cf968fa560a9e3cf9b6c84516032fca8 /gc/doc/gcdescr.html | |
| parent | Updates from 0.2.1 into 0.2.1-inu-1.5 (diff) | |
| download | w3m-85da7ee692072c643939e9f4b24fbd1e74e64e70.tar.gz w3m-85da7ee692072c643939e9f4b24fbd1e74e64e70.zip | |
Update to w3m-0.2.1-inu-1.6.
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diff --git a/gc/doc/gcdescr.html b/gc/doc/gcdescr.html new file mode 100644 index 0000000..65e8a8f --- /dev/null +++ b/gc/doc/gcdescr.html @@ -0,0 +1,438 @@ +<HTML> +<HEAD> + <TITLE> Conservative GC Algorithmic Overview </TITLE> + <AUTHOR> Hans-J. Boehm, Silicon Graphics</author> +</HEAD> +<BODY> +<H1> <I>This is under construction</i> </h1> +<H1> Conservative GC Algorithmic Overview </h1> +<P> +This is a description of the algorithms and data structures used in our +conservative garbage collector. I expect the level of detail to increase +with time. For a survey of GC algorithms, see for example +<A HREF="ftp://ftp.cs.utexas.edu/pub/garbage/gcsurvey.ps"> Paul Wilson's +excellent paper</a>. For an overview of the collector interface, +see <A HREF="gcinterface.html">here</a>. +<P> +This description is targeted primarily at someone trying to understand the +source code. It specifically refers to variable and function names. +It may also be useful for understanding the algorithms at a higher level. +<P> +The description here assumes that the collector is used in default mode. +In particular, we assume that it used as a garbage collector, and not just +a leak detector. We initially assume that it is used in stop-the-world, +non-incremental mode, though the presence of the incremental collector +will be apparent in the design. +We assume the default finalization model, but the code affected by that +is very localized. +<H2> Introduction </h2> +The garbage collector uses a modified mark-sweep algorithm. Conceptually +it operates roughly in four phases: + +<OL> + +<LI> +<I>Preparation</i> Clear all mark bits, indicating that all objects +are potentially unreachable. + +<LI> +<I>Mark phase</i> Marks all objects that can be reachable via chains of +pointers from variables. Normally the collector has no real information +about the location of pointer variables in the heap, so it +views all static data areas, stacks and registers as potentially containing +containing pointers. Any bit patterns that represent addresses inside +heap objects managed by the collector are viewed as pointers. +Unless the client program has made heap object layout information +available to the collector, any heap objects found to be reachable from +variables are again scanned similarly. + +<LI> +<I>Sweep phase</i> Scans the heap for inaccessible, and hence unmarked, +objects, and returns them to an appropriate free list for reuse. This is +not really a separate phase; even in non incremental mode this is operation +is usually performed on demand during an allocation that discovers an empty +free list. Thus the sweep phase is very unlikely to touch a page that +would not have been touched shortly thereafter anyway. + +<LI> +<I>Finalization phase</i> Unreachable objects which had been registered +for finalization are enqueued for finalization outside the collector. + +</ol> + +<P> +The remaining sections describe the memory allocation data structures, +and then the last 3 collection phases in more detail. We conclude by +outlining some of the additional features implemented in the collector. + +<H2>Allocation</h2> +The collector includes its own memory allocator. The allocator obtains +memory from the system in a platform-dependent way. Under UNIX, it +uses either <TT>malloc</tt>, <TT>sbrk</tt>, or <TT>mmap</tt>. +<P> +Most static data used by the allocator, as well as that needed by the +rest of the garbage collector is stored inside the +<TT>_GC_arrays</tt> structure. +This allows the garbage collector to easily ignore the collectors own +data structures when it searches for root pointers. Other allocator +and collector internal data structures are allocated dynamically +with <TT>GC_scratch_alloc</tt>. <TT>GC_scratch_alloc</tt> does not +allow for deallocation, and is therefore used only for permanent data +structures. +<P> +The allocator allocates objects of different <I>kinds</i>. +Different kinds are handled somewhat differently by certain parts +of the garbage collector. Certain kinds are scanned for pointers, +others are not. Some may have per-object type descriptors that +determine pointer locations. Or a specific kind may correspond +to one specific object layout. Two built-in kinds are uncollectable. +One (<TT>STUBBORN</tt>) is immutable without special precautions. +In spite of that, it is very likely that most applications currently +use at most two kinds: <TT>NORMAL</tt> and <TT>PTRFREE</tt> objects. +<P> +The collector uses a two level allocator. A large block is defined to +be one larger than half of <TT>HBLKSIZE</tt>, which is a power of 2, +typically on the order of the page size. +<P> +Large block sizes are rounded up to +the next multiple of <TT>HBLKSIZE</tt> and then allocated by +<TT>GC_allochblk</tt>. This uses roughly what Paul Wilson has termed +a "next fit" algorithm, i.e. first-fit with a rotating pointer. +The implementation does check for a better fitting immediately +adjacent block, which gives it somewhat better fragmentation characteristics. +I'm now convinced it should use a best fit algorithm. The actual +implementation of <TT>GC_allochblk</tt> +is significantly complicated by black-listing issues +(see below). +<P> +Small blocks are allocated in blocks of size <TT>HBLKSIZE</tt>. +Each block is +dedicated to only one object size and kind. The allocator maintains +separate free lists for each size and kind of object. +<P> +In order to avoid allocating blocks for too many distinct object sizes, +the collector normally does not directly allocate objects of every possible +request size. Instead request are rounded up to one of a smaller number +of allocated sizes, for which free lists are maintained. The exact +allocated sizes are computed on demand, but subject to the constraint +that they increase roughly in geometric progression. Thus objects +requested early in the execution are likely to be allocated with exactly +the requested size, subject to alignment constraints. +See <TT>GC_init_size_map</tt> for details. +<P> +The actual size rounding operation during small object allocation is +implemented as a table lookup in <TT>GC_size_map</tt>. +<P> +Both collector initialization and computation of allocated sizes are +handled carefully so that they do not slow down the small object fast +allocation path. An attempt to allocate before the collector is initialized, +or before the appropriate <TT>GC_size_map</tt> entry is computed, +will take the same path as an allocation attempt with an empty free list. +This results in a call to the slow path code (<TT>GC_generic_malloc_inner</tt>) +which performs the appropriate initialization checks. +<P> +In non-incremental mode, we make a decision about whether to garbage collect +whenever an allocation would otherwise have failed with the current heap size. +If the total amount of allocation since the last collection is less than +the heap size divided by <TT>GC_free_space_divisor</tt>, we try to +expand the heap. Otherwise, we initiate a garbage collection. This ensures +that the amount of garbage collection work per allocated byte remains +constant. +<P> +The above is in fat an oversimplification of the real heap expansion +heuristic, which adjusts slightly for root size and certain kinds of +fragmentation. In particular, programs with a large root set size and +little live heap memory will expand the heap to amortize the cost of +scanning the roots. +<P> +Versions 5.x of the collector actually collect more frequently in +nonincremental mode. The large block allocator usually refuses to split +large heap blocks once the garbage collection threshold is +reached. This often has the effect of collecting well before the +heap fills up, thus reducing fragmentation and working set size at the +expense of GC time. 6.x will chose an intermediate strategy depending +on how much large object allocation has taken place in the past. +(If the collector is configured to unmap unused pages, versions 6.x +will use the 5.x strategy.) +<P> +(It has been suggested that this should be adjusted so that we favor +expansion if the resulting heap still fits into physical memory. +In many cases, that would no doubt help. But it is tricky to do this +in a way that remains robust if multiple application are contending +for a single pool of physical memory.) + +<H2>Mark phase</h2> + +The marker maintains an explicit stack of memory regions that are known +to be accessible, but that have not yet been searched for contained pointers. +Each stack entry contains the starting address of the block to be scanned, +as well as a descriptor of the block. If no layout information is +available for the block, then the descriptor is simply a length. +(For other possibilities, see <TT>gc_mark.h</tt>.) +<P> +At the beginning of the mark phase, all root segments are pushed on the +stack by <TT>GC_push_roots</tt>. If <TT>ALL_INTERIOR_PTRS</tt> is not +defined, then stack roots require special treatment. In this case, the +normal marking code ignores interior pointers, but <TT>GC_push_all_stack</tt> +explicitly checks for interior pointers and pushes descriptors for target +objects. +<P> +The marker is structured to allow incremental marking. +Each call to <TT>GC_mark_some</tt> performs a small amount of +work towards marking the heap. +It maintains +explicit state in the form of <TT>GC_mark_state</tt>, which +identifies a particular sub-phase. Some other pieces of state, most +notably the mark stack, identify how much work remains to be done +in each sub-phase. The normal progression of mark states for +a stop-the-world collection is: +<OL> +<LI> <TT>MS_INVALID</tt> indicating that there may be accessible unmarked +objects. In this case <TT>GC_objects_are_marked</tt> will simultaneously +be false, so the mark state is advanced to +<LI> <TT>MS_PUSH_UNCOLLECTABLE</tt> indicating that it suffices to push +uncollectable objects, roots, and then mark everything reachable from them. +<TT>Scan_ptr</tt> is advanced through the heap until all uncollectable +objects are pushed, and objects reachable from them are marked. +At that point, the next call to <TT>GC_mark_some</tt> calls +<TT>GC_push_roots</tt> to push the roots. It the advances the +mark state to +<LI> <TT>MS_ROOTS_PUSHED</tt> asserting that once the mark stack is +empty, all reachable objects are marked. Once in this state, we work +only on emptying the mark stack. Once this is completed, the state +changes to +<LI> <TT>MS_NONE</tt> indicating that reachable objects are marked. +</ol> + +The core mark routine <TT>GC_mark_from_mark_stack</tt>, is called +repeatedly by several of the sub-phases when the mark stack starts to fill +up. It is also called repeatedly in <TT>MS_ROOTS_PUSHED</tt> state +to empty the mark stack. +The routine is designed to only perform a limited amount of marking at +each call, so that it can also be used by the incremental collector. +It is fairly carefully tuned, since it usually consumes a large majority +of the garbage collection time. +<P> +The marker correctly handles mark stack overflows. Whenever the mark stack +overflows, the mark state is reset to <TT>MS_INVALID</tt>. +Since there are already marked objects in the heap, +this eventually forces a complete +scan of the heap, searching for pointers, during which any unmarked objects +referenced by marked objects are again pushed on the mark stack. This +process is repeated until the mark phase completes without a stack overflow. +Each time the stack overflows, an attempt is made to grow the mark stack. +All pieces of the collector that push regions onto the mark stack have to be +careful to ensure forward progress, even in case of repeated mark stack +overflows. Every mark attempt results in additional marked objects. +<P> +Each mark stack entry is processed by examining all candidate pointers +in the range described by the entry. If the region has no associated +type information, then this typically requires that each 4-byte aligned +quantity (8-byte aligned with 64-bit pointers) be considered a candidate +pointer. +<P> +We determine whether a candidate pointer is actually the address of +a heap block. This is done in the following steps: +<NL> +<LI> The candidate pointer is checked against rough heap bounds. +These heap bounds are maintained such that all actual heap objects +fall between them. In order to facilitate black-listing (see below) +we also include address regions that the heap is likely to expand into. +Most non-pointers fail this initial test. +<LI> The candidate pointer is divided into two pieces; the most significant +bits identify a <TT>HBLKSIZE</tt>-sized page in the address space, and +the least significant bits specify an offset within that page. +(A hardware page may actually consist of multiple such pages. +HBLKSIZE is usually the page size divided by a small power of two.) +<LI> +The page address part of the candidate pointer is looked up in a +<A HREF="tree.html">table</a>. +Each table entry contains either 0, indicating that the page is not part +of the garbage collected heap, a small integer <I>n</i>, indicating +that the page is part of large object, starting at least <I>n</i> pages +back, or a pointer to a descriptor for the page. In the first case, +the candidate pointer i not a true pointer and can be safely ignored. +In the last two cases, we can obtain a descriptor for the page containing +the beginning of the object. +<LI> +The starting address of the referenced object is computed. +The page descriptor contains the size of the object(s) +in that page, the object kind, and the necessary mark bits for those +objects. The size information can be used to map the candidate pointer +to the object starting address. To accelerate this process, the page header +also contains a pointer to a precomputed map of page offsets to displacements +from the beginning of an object. The use of this map avoids a +potentially slow integer remainder operation in computing the object +start address. +<LI> +The mark bit for the target object is checked and set. If the object +was previously unmarked, the object is pushed on the mark stack. +The descriptor is read from the page descriptor. (This is computed +from information <TT>GC_obj_kinds</tt> when the page is first allocated.) +</nl> +<P> +At the end of the mark phase, mark bits for left-over free lists are cleared, +in case a free list was accidentally marked due to a stray pointer. + +<H2>Sweep phase</h2> + +At the end of the mark phase, all blocks in the heap are examined. +Unmarked large objects are immediately returned to the large object free list. +Each small object page is checked to see if all mark bits are clear. +If so, the entire page is returned to the large object free list. +Small object pages containing some reachable object are queued for later +sweeping. +<P> +This initial sweep pass touches only block headers, not +the blocks themselves. Thus it does not require significant paging, even +if large sections of the heap are not in physical memory. +<P> +Nonempty small object pages are swept when an allocation attempt +encounters an empty free list for that object size and kind. +Pages for the correct size and kind are repeatedly swept until at +least one empty block is found. Sweeping such a page involves +scanning the mark bit array in the page header, and building a free +list linked through the first words in the objects themselves. +This does involve touching the appropriate data page, but in most cases +it will be touched only just before it is used for allocation. +Hence any paging is essentially unavoidable. +<P> +Except in the case of pointer-free objects, we maintain the invariant +that any object in a small object free list is cleared (except possibly +for the link field). Thus it becomes the burden of the small object +sweep routine to clear objects. This has the advantage that we can +easily recover from accidentally marking a free list, though that could +also be handled by other means. The collector currently spends a fair +amount of time clearing objects, and this approach should probably be +revisited. +<P> +In most configurations, we use specialized sweep routines to handle common +small object sizes. Since we allocate one mark bit per word, it becomes +easier to examine the relevant mark bits if the object size divides +the word length evenly. We also suitably unroll the inner sweep loop +in each case. (It is conceivable that profile-based procedure cloning +in the compiler could make this unnecessary and counterproductive. I +know of no existing compiler to which this applies.) +<P> +The sweeping of small object pages could be avoided completely at the expense +of examining mark bits directly in the allocator. This would probably +be more expensive, since each allocation call would have to reload +a large amount of state (e.g. next object address to be swept, position +in mark bit table) before it could do its work. The current scheme +keeps the allocator simple and allows useful optimizations in the sweeper. + +<H2>Finalization</h2> +Both <TT>GC_register_disappearing_link</tt> and +<TT>GC_register_finalizer</tt> add the request to a corresponding hash +table. The hash table is allocated out of collected memory, but +the reference to the finalizable object is hidden from the collector. +Currently finalization requests are processed non-incrementally at the +end of a mark cycle. +<P> +The collector makes an initial pass over the table of finalizable objects, +pushing the contents of unmarked objects onto the mark stack. +After pushing each object, the marker is invoked to mark all objects +reachable from it. The object itself is not explicitly marked. +This assures that objects on which a finalizer depends are neither +collected nor finalized. +<P> +If in the process of marking from an object the +object itself becomes marked, we have uncovered +a cycle involving the object. This usually results in a warning from the +collector. Such objects are not finalized, since it may be +unsafe to do so. See the more detailed +<A HREF="finalization.html"> discussion of finalization semantics</a>. +<P> +Any objects remaining unmarked at the end of this process are added to +a queue of objects whose finalizers can be run. Depending on collector +configuration, finalizers are dequeued and run either implicitly during +allocation calls, or explicitly in response to a user request. +<P> +The collector provides a mechanism for replacing the procedure that is +used to mark through objects. This is used both to provide support for +Java-style unordered finalization, and to ignore certain kinds of cycles, +<I>e.g.</i> those arising from C++ implementations of virtual inheritance. + +<H2>Generational Collection and Dirty Bits</h2> +We basically use the parallel and generational GC algorithm described in +<A HREF="papers/pldi91.ps.gz">"Mostly Parallel Garbage Collection"</a>, +by Boehm, Demers, and Shenker. +<P> +The most significant modification is that +the collector always runs in the allocating thread. +There is no separate garbage collector thread. +If an allocation attempt either requests a large object, or encounters +an empty small object free list, and notices that there is a collection +in progress, it immediately performs a small amount of marking work +as described above. +<P> +This change was made both because we wanted to easily accommodate +single-threaded environments, and because a separate GC thread requires +very careful control over the scheduler to prevent the mutator from +out-running the collector, and hence provoking unneeded heap growth. +<P> +In incremental mode, the heap is always expanded when we encounter +insufficient space for an allocation. Garbage collection is triggered +whenever we notice that more than +<TT>GC_heap_size</tt>/2 * <TT>GC_free_space_divisor</tt> +bytes of allocation have taken place. +After <TT>GC_full_freq</tt> minor collections a major collection +is started. +<P> +All collections initially run interrupted until a predetermined +amount of time (50 msecs by default) has expired. If this allows +the collection to complete entirely, we can avoid correcting +for data structure modifications during the collection. If it does +not complete, we return control to the mutator, and perform small +amounts of additional GC work during those later allocations that +cannot be satisfied from small object free lists. When marking completes, +the set of modified pages is retrieved, and we mark once again from +marked objects on those pages, this time with the mutator stopped. +<P> +We keep track of modified pages using one of three distinct mechanisms: +<OL> +<LI> +Through explicit mutator cooperation. Currently this requires +the use of <TT>GC_malloc_stubborn</tt>. +<LI> +By write-protecting physical pages and catching write faults. This is +implemented for many Unix-like systems and for win32. It is not possible +in a few environments. +<LI> +By retrieving dirty bit information from /proc. (Currently only Sun's +Solaris supports this. Though this is considerably cleaner, performance +may actually be better with mprotect and signals.) +</ol> + +<H2>Thread support</h2> +We support several different threading models. Unfortunately Pthreads, +the only reasonably well standardized thread model, supports too narrow +an interface for conservative garbage collection. There appears to be +no portable way to allow the collector to coexist with various Pthreads +implementations. Hence we currently support only a few of the more +common Pthreads implementations. +<P> +In particular, it is very difficult for the collector to stop all other +threads in the system and examine the register contents. This is currently +accomplished with very different mechanisms for different Pthreads +implementations. The Solaris implementation temporarily disables much +of the user-level threads implementation by stopping kernel-level threads +("lwp"s). The Irix implementation sends signals to individual Pthreads +and has them wait in the signal handler. The Linux implementation +is similar in spirit to the Irix one. +<P> +The Irix implementation uses +only documented Pthreads calls, but relies on extensions to their semantics, +notably the use of mutexes and condition variables from signal +handlers. The Linux implementation should be far closer to +portable, though impirically it is not completely portable. +<P> +All implementations must +intercept thread creation and a few other thread-specific calls to allow +enumeration of threads and location of thread stacks. This is current +accomplished with <TT># define</tt>'s in <TT>gc.h</tt>, or optionally +by using ld's function call wrapping mechanism under Linux. +<P> +Comments are appreciated. Please send mail to +<A HREF="mailto:boehm@acm.org"><TT>boehm@acm.org</tt></a> +</body> |