439 lines
22 KiB
HTML
439 lines
22 KiB
HTML
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<HTML>
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<HEAD>
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<TITLE> Conservative GC Algorithmic Overview </TITLE>
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<AUTHOR> Hans-J. Boehm, Silicon Graphics</author>
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</HEAD>
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<BODY>
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<H1> <I>This is under construction</i> </h1>
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<H1> Conservative GC Algorithmic Overview </h1>
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<P>
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This is a description of the algorithms and data structures used in our
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conservative garbage collector. I expect the level of detail to increase
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with time. For a survey of GC algorithms, see for example
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<A HREF="ftp://ftp.cs.utexas.edu/pub/garbage/gcsurvey.ps"> Paul Wilson's
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excellent paper</a>. For an overview of the collector interface,
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see <A HREF="gcinterface.html">here</a>.
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<P>
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This description is targeted primarily at someone trying to understand the
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source code. It specifically refers to variable and function names.
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It may also be useful for understanding the algorithms at a higher level.
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<P>
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The description here assumes that the collector is used in default mode.
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In particular, we assume that it used as a garbage collector, and not just
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a leak detector. We initially assume that it is used in stop-the-world,
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non-incremental mode, though the presence of the incremental collector
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will be apparent in the design.
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We assume the default finalization model, but the code affected by that
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is very localized.
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<H2> Introduction </h2>
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The garbage collector uses a modified mark-sweep algorithm. Conceptually
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it operates roughly in four phases:
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<OL>
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<LI>
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<I>Preparation</i> Clear all mark bits, indicating that all objects
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are potentially unreachable.
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<LI>
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<I>Mark phase</i> Marks all objects that can be reachable via chains of
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pointers from variables. Normally the collector has no real information
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about the location of pointer variables in the heap, so it
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views all static data areas, stacks and registers as potentially containing
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containing pointers. Any bit patterns that represent addresses inside
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heap objects managed by the collector are viewed as pointers.
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Unless the client program has made heap object layout information
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available to the collector, any heap objects found to be reachable from
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variables are again scanned similarly.
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<LI>
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<I>Sweep phase</i> Scans the heap for inaccessible, and hence unmarked,
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objects, and returns them to an appropriate free list for reuse. This is
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not really a separate phase; even in non incremental mode this is operation
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is usually performed on demand during an allocation that discovers an empty
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free list. Thus the sweep phase is very unlikely to touch a page that
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would not have been touched shortly thereafter anyway.
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<LI>
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<I>Finalization phase</i> Unreachable objects which had been registered
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for finalization are enqueued for finalization outside the collector.
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</ol>
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<P>
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The remaining sections describe the memory allocation data structures,
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and then the last 3 collection phases in more detail. We conclude by
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outlining some of the additional features implemented in the collector.
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<H2>Allocation</h2>
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The collector includes its own memory allocator. The allocator obtains
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memory from the system in a platform-dependent way. Under UNIX, it
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uses either <TT>malloc</tt>, <TT>sbrk</tt>, or <TT>mmap</tt>.
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<P>
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Most static data used by the allocator, as well as that needed by the
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rest of the garbage collector is stored inside the
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<TT>_GC_arrays</tt> structure.
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This allows the garbage collector to easily ignore the collectors own
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data structures when it searches for root pointers. Other allocator
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and collector internal data structures are allocated dynamically
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with <TT>GC_scratch_alloc</tt>. <TT>GC_scratch_alloc</tt> does not
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allow for deallocation, and is therefore used only for permanent data
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structures.
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<P>
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The allocator allocates objects of different <I>kinds</i>.
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Different kinds are handled somewhat differently by certain parts
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of the garbage collector. Certain kinds are scanned for pointers,
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others are not. Some may have per-object type descriptors that
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determine pointer locations. Or a specific kind may correspond
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to one specific object layout. Two built-in kinds are uncollectable.
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One (<TT>STUBBORN</tt>) is immutable without special precautions.
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In spite of that, it is very likely that most applications currently
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use at most two kinds: <TT>NORMAL</tt> and <TT>PTRFREE</tt> objects.
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<P>
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The collector uses a two level allocator. A large block is defined to
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be one larger than half of <TT>HBLKSIZE</tt>, which is a power of 2,
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typically on the order of the page size.
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<P>
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Large block sizes are rounded up to
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the next multiple of <TT>HBLKSIZE</tt> and then allocated by
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<TT>GC_allochblk</tt>. This uses roughly what Paul Wilson has termed
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a "next fit" algorithm, i.e. first-fit with a rotating pointer.
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The implementation does check for a better fitting immediately
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adjacent block, which gives it somewhat better fragmentation characteristics.
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I'm now convinced it should use a best fit algorithm. The actual
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implementation of <TT>GC_allochblk</tt>
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is significantly complicated by black-listing issues
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(see below).
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<P>
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Small blocks are allocated in blocks of size <TT>HBLKSIZE</tt>.
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Each block is
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dedicated to only one object size and kind. The allocator maintains
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separate free lists for each size and kind of object.
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<P>
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In order to avoid allocating blocks for too many distinct object sizes,
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the collector normally does not directly allocate objects of every possible
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request size. Instead request are rounded up to one of a smaller number
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of allocated sizes, for which free lists are maintained. The exact
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allocated sizes are computed on demand, but subject to the constraint
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that they increase roughly in geometric progression. Thus objects
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requested early in the execution are likely to be allocated with exactly
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the requested size, subject to alignment constraints.
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See <TT>GC_init_size_map</tt> for details.
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<P>
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The actual size rounding operation during small object allocation is
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implemented as a table lookup in <TT>GC_size_map</tt>.
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<P>
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Both collector initialization and computation of allocated sizes are
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handled carefully so that they do not slow down the small object fast
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allocation path. An attempt to allocate before the collector is initialized,
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or before the appropriate <TT>GC_size_map</tt> entry is computed,
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will take the same path as an allocation attempt with an empty free list.
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This results in a call to the slow path code (<TT>GC_generic_malloc_inner</tt>)
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which performs the appropriate initialization checks.
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<P>
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In non-incremental mode, we make a decision about whether to garbage collect
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whenever an allocation would otherwise have failed with the current heap size.
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If the total amount of allocation since the last collection is less than
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the heap size divided by <TT>GC_free_space_divisor</tt>, we try to
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expand the heap. Otherwise, we initiate a garbage collection. This ensures
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that the amount of garbage collection work per allocated byte remains
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constant.
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<P>
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The above is in fat an oversimplification of the real heap expansion
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heuristic, which adjusts slightly for root size and certain kinds of
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fragmentation. In particular, programs with a large root set size and
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little live heap memory will expand the heap to amortize the cost of
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scanning the roots.
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<P>
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Versions 5.x of the collector actually collect more frequently in
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nonincremental mode. The large block allocator usually refuses to split
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large heap blocks once the garbage collection threshold is
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reached. This often has the effect of collecting well before the
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heap fills up, thus reducing fragmentation and working set size at the
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expense of GC time. 6.x will chose an intermediate strategy depending
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on how much large object allocation has taken place in the past.
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(If the collector is configured to unmap unused pages, versions 6.x
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will use the 5.x strategy.)
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<P>
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(It has been suggested that this should be adjusted so that we favor
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expansion if the resulting heap still fits into physical memory.
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In many cases, that would no doubt help. But it is tricky to do this
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in a way that remains robust if multiple application are contending
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for a single pool of physical memory.)
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<H2>Mark phase</h2>
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The marker maintains an explicit stack of memory regions that are known
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to be accessible, but that have not yet been searched for contained pointers.
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Each stack entry contains the starting address of the block to be scanned,
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as well as a descriptor of the block. If no layout information is
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available for the block, then the descriptor is simply a length.
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(For other possibilities, see <TT>gc_mark.h</tt>.)
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<P>
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At the beginning of the mark phase, all root segments are pushed on the
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stack by <TT>GC_push_roots</tt>. If <TT>ALL_INTERIOR_PTRS</tt> is not
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defined, then stack roots require special treatment. In this case, the
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normal marking code ignores interior pointers, but <TT>GC_push_all_stack</tt>
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explicitly checks for interior pointers and pushes descriptors for target
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objects.
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<P>
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The marker is structured to allow incremental marking.
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Each call to <TT>GC_mark_some</tt> performs a small amount of
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work towards marking the heap.
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It maintains
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explicit state in the form of <TT>GC_mark_state</tt>, which
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identifies a particular sub-phase. Some other pieces of state, most
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notably the mark stack, identify how much work remains to be done
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in each sub-phase. The normal progression of mark states for
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a stop-the-world collection is:
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<OL>
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<LI> <TT>MS_INVALID</tt> indicating that there may be accessible unmarked
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objects. In this case <TT>GC_objects_are_marked</tt> will simultaneously
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be false, so the mark state is advanced to
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<LI> <TT>MS_PUSH_UNCOLLECTABLE</tt> indicating that it suffices to push
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uncollectable objects, roots, and then mark everything reachable from them.
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<TT>Scan_ptr</tt> is advanced through the heap until all uncollectable
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objects are pushed, and objects reachable from them are marked.
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At that point, the next call to <TT>GC_mark_some</tt> calls
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<TT>GC_push_roots</tt> to push the roots. It the advances the
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mark state to
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<LI> <TT>MS_ROOTS_PUSHED</tt> asserting that once the mark stack is
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empty, all reachable objects are marked. Once in this state, we work
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only on emptying the mark stack. Once this is completed, the state
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changes to
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<LI> <TT>MS_NONE</tt> indicating that reachable objects are marked.
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</ol>
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The core mark routine <TT>GC_mark_from_mark_stack</tt>, is called
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repeatedly by several of the sub-phases when the mark stack starts to fill
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up. It is also called repeatedly in <TT>MS_ROOTS_PUSHED</tt> state
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to empty the mark stack.
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The routine is designed to only perform a limited amount of marking at
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each call, so that it can also be used by the incremental collector.
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It is fairly carefully tuned, since it usually consumes a large majority
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of the garbage collection time.
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<P>
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The marker correctly handles mark stack overflows. Whenever the mark stack
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overflows, the mark state is reset to <TT>MS_INVALID</tt>.
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Since there are already marked objects in the heap,
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this eventually forces a complete
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scan of the heap, searching for pointers, during which any unmarked objects
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referenced by marked objects are again pushed on the mark stack. This
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process is repeated until the mark phase completes without a stack overflow.
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Each time the stack overflows, an attempt is made to grow the mark stack.
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All pieces of the collector that push regions onto the mark stack have to be
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careful to ensure forward progress, even in case of repeated mark stack
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overflows. Every mark attempt results in additional marked objects.
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<P>
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Each mark stack entry is processed by examining all candidate pointers
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in the range described by the entry. If the region has no associated
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type information, then this typically requires that each 4-byte aligned
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quantity (8-byte aligned with 64-bit pointers) be considered a candidate
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pointer.
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<P>
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We determine whether a candidate pointer is actually the address of
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a heap block. This is done in the following steps:
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<NL>
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<LI> The candidate pointer is checked against rough heap bounds.
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These heap bounds are maintained such that all actual heap objects
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fall between them. In order to facilitate black-listing (see below)
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we also include address regions that the heap is likely to expand into.
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Most non-pointers fail this initial test.
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<LI> The candidate pointer is divided into two pieces; the most significant
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bits identify a <TT>HBLKSIZE</tt>-sized page in the address space, and
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the least significant bits specify an offset within that page.
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(A hardware page may actually consist of multiple such pages.
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HBLKSIZE is usually the page size divided by a small power of two.)
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<LI>
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The page address part of the candidate pointer is looked up in a
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<A HREF="tree.html">table</a>.
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Each table entry contains either 0, indicating that the page is not part
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of the garbage collected heap, a small integer <I>n</i>, indicating
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that the page is part of large object, starting at least <I>n</i> pages
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back, or a pointer to a descriptor for the page. In the first case,
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the candidate pointer i not a true pointer and can be safely ignored.
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In the last two cases, we can obtain a descriptor for the page containing
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the beginning of the object.
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<LI>
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The starting address of the referenced object is computed.
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The page descriptor contains the size of the object(s)
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in that page, the object kind, and the necessary mark bits for those
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objects. The size information can be used to map the candidate pointer
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to the object starting address. To accelerate this process, the page header
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also contains a pointer to a precomputed map of page offsets to displacements
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from the beginning of an object. The use of this map avoids a
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potentially slow integer remainder operation in computing the object
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start address.
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<LI>
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The mark bit for the target object is checked and set. If the object
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was previously unmarked, the object is pushed on the mark stack.
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The descriptor is read from the page descriptor. (This is computed
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from information <TT>GC_obj_kinds</tt> when the page is first allocated.)
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</nl>
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<P>
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At the end of the mark phase, mark bits for left-over free lists are cleared,
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in case a free list was accidentally marked due to a stray pointer.
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<H2>Sweep phase</h2>
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At the end of the mark phase, all blocks in the heap are examined.
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Unmarked large objects are immediately returned to the large object free list.
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Each small object page is checked to see if all mark bits are clear.
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If so, the entire page is returned to the large object free list.
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Small object pages containing some reachable object are queued for later
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sweeping.
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<P>
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This initial sweep pass touches only block headers, not
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the blocks themselves. Thus it does not require significant paging, even
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if large sections of the heap are not in physical memory.
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<P>
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Nonempty small object pages are swept when an allocation attempt
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encounters an empty free list for that object size and kind.
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Pages for the correct size and kind are repeatedly swept until at
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least one empty block is found. Sweeping such a page involves
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scanning the mark bit array in the page header, and building a free
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list linked through the first words in the objects themselves.
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This does involve touching the appropriate data page, but in most cases
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it will be touched only just before it is used for allocation.
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Hence any paging is essentially unavoidable.
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<P>
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Except in the case of pointer-free objects, we maintain the invariant
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that any object in a small object free list is cleared (except possibly
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for the link field). Thus it becomes the burden of the small object
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sweep routine to clear objects. This has the advantage that we can
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easily recover from accidentally marking a free list, though that could
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also be handled by other means. The collector currently spends a fair
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amount of time clearing objects, and this approach should probably be
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revisited.
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<P>
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In most configurations, we use specialized sweep routines to handle common
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small object sizes. Since we allocate one mark bit per word, it becomes
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easier to examine the relevant mark bits if the object size divides
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the word length evenly. We also suitably unroll the inner sweep loop
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in each case. (It is conceivable that profile-based procedure cloning
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in the compiler could make this unnecessary and counterproductive. I
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know of no existing compiler to which this applies.)
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<P>
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The sweeping of small object pages could be avoided completely at the expense
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of examining mark bits directly in the allocator. This would probably
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be more expensive, since each allocation call would have to reload
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a large amount of state (e.g. next object address to be swept, position
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in mark bit table) before it could do its work. The current scheme
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keeps the allocator simple and allows useful optimizations in the sweeper.
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|
|
|
||
|
|
<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>
|