Way guard: a segmented counting bloom filter approach to reducing energy for set-associative caches.
Way Guard: A Segmented Counting Bloom Filter Approach
to Reducing Energy for Set-Associative Caches
Hsien-Hsin S. Lee⋆
†ARM Inc., Austin, TX
‡ARM Ltd., Cambridge, UK
⋆Georgia Tech, Atlanta, GA
The design trend of caches in modern processors continues to in-
crease their capacity with higher associativity to cope with large
data footprint and take advantage of feature size shrink, which, un-
fortunately, also leads to higher energy consumption. This paper
presents a technique using segmented counting Bloom filters called
“Way Guard” to reduce the number of redundant way lookups in
large set-associative caches to achieve dynamic energy savings.
Our Way Guard mechanism only looks up an average of 25-30%
of the cache ways and saved up to 65% of the L2 energy and up to
70% of the L1 cache energy.
Categories and Subject Descriptors
C.1.0 [Processor Architecture]: General; B.3.2 [Memory Struc-
tures]: Design Styles– Cache memories
Cache hierarchy has become a main consumer of both static and
dynamic energy in processors. Even so, the trend in modern pro-
cessor designs continues to increase both capacity and associativ-
ity to accommodate the ever-growing workloads and alleviate con-
flict misses. For processors employing highly associative caches,
the energy consumption gets even worse as N-tag comparisons are
needed for each parallel lookup of an N-way cache. In fact, most of
theenergy consumed for suchlookups isredundant as therequested
data can only be present in one particular way. This redundancy
provides a good opportunity for saving dynamic energy.
In this paper, we propose a technique called Way Guard based
on segmented counting Bloom filters to exploit these energy sav-
ing opportunities. Our scheme uses counting Bloom filters to ef-
ficiently skip the lookup of cache lines that do not contain the re-
quested data to save significant energy in cache accesses. Bloom
filters are simple, fast structures that can eliminate the need of
performing associative lookup especially when the lookup address
space is huge. They can replace the expensive set-associative tag
matching with a simple bit vector that precisely identifies addresses
that have not been observed before. This mechanism provides early
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L−bit Counters Bit Vector
(b) Counting Bloom Filter
Figure 1: Bloom Filters
detection of events to avoid an associative buffer lookup. This im-
proves energy consumption significantly without adversely affect-
ing performance given the efficient hardware structures.
The rest of this paper is arranged as follows. Section 2 explains
Bloom filters. Section 3 describes our segmented Bloom filter and
its energy-saving feature. Section 4 explains Way Guard technique.
Section 5 describes our simulation methodology and analysis. Sec-
tion 6 reviews prior techniques. Section 7 concludes.
The original Bloom filter concept is depicted in Figure 1(a). It
consists of several hash functions and a bit vector. A given N-bit
address is hashed into k hash values using k different random hash
functions. The output of each hash function is an m-bit index that
addresses the 2mentry bit vector, where m is much smaller than N.
Initially, the Bloom filter bit vector is zero. Whenever an N-
bit address is observed, it is hashed to the bit vector and the bit
value hashed by each m-bit index is set to one. When a query is
to be made, the given N-bit address is hashed using the same hash
functions and the bit values are read from the locations indexed by
the m-bit hash value. If at least one of the bits is 0, it indicates that
this address was definitely not observed before. This is called atrue
miss. Whereas, if all of the bit values are 1, the address may have
been observed but with no guarantee, which is called a false hit.
As the number of hash functions increases, the Bloom filter bit
vector is polluted faster. On the other hand, the probability of find-
ing a zero during a query is increased if more hash functions are
used. The major drawback of the original Bloom filter is the high
false hit rate as it can be quickly filled up with all 1’s. Also, once a
bit is set, there is no way to reset it. Thus, as more bits are set, the
number of false hits increases. To address this issue, the counting
Bloom filter  shown in Figure 1(b) was proposed for web cache
sharing to provide capability of resetting entries in the filter. First,
an array of counters is added along with the bit vector of the orig-
inal Bloom Filter. Each L-bit counter has a one-to-one association
with each bit in the bit vector. Queries to a counting Bloom filter
is similar with a slight modification: when an address is entered,
2 − 1
Figure 2: Segmented Bloom filter
each m-bit hash index will increment its corresponding counter of
the counter array in addition to setting the bit vector. Similarly,
when an address is removed from the Bloom filter, each m-bit hash
index will decrement its corresponding counter. If more than one
hash indexes to the same location for a given address, the counter is
incremented or decremented only once. Finally, when a counter is
reduced to zero, its associated bit in the bit vector will be cleared.
3. SEGMENTED COUNTINGBLOOMFIL-
One application of the counting Bloom filter is to keep track of
the line-fills and replacements of a cache and indicate whether an
address is present in the cache. Query to a counting Bloom filter
consumes less energy and quicker than accessing the entire cache.
Ghosh et al. has shown a cache miss detection technique using
a segmented counting Bloom filter . Their design redrawn in
Figure 2 contains the counter array (L bits per counter) decoupled
from the bit vector with a duplicated hash function on the bit vector
side. The cache line fill/eviction addresses are sent to the counter
array using one hash function while the cache request address from
the processor is sent to the bit vector using a copy of the same hash
function. The segmented Bloom filter design allows the counter
array and bit vector to operate in separate physical locations.
There are several reasons for a segmented Bloom filter: 1) We
only need the bit vector, which is smaller than the counter, to obtain
the outcome of a query. Decoupling the bit vector enables faster
and lower energy accesses to the Bloom Filter. Hence the result of
a query issued from the core can be obtained by just looking up the
bit vector; 2) The update to the counters is not time-critical with
respect to the core. So, the segmented design allows the counter
array to run at lower frequency than the bit vector. The vector part
being smaller provides fast access time, whereas the larger counter
part runs at a lower frequency to save energy. The only additional
overhead of the segmented design is the duplication of the hash
function hardware. We now describe an innovative application of
the segmented counting Bloom filter to avoid unnecessary cache
4. WAY GUARD MECHANISM
Inthis section, wedescribe anovel applicationof countingBloom
filters to set-associative caches to determine data presence and save
lookup energy. Figure 3 illustrates the design of a 4-way cache
with our proposed Way Guard mechanism. As shown, each Way
Guard, structurally the same as the counting Bloom Filter, consists
of a Bit Vector (shown as BV) and an array of counters. Each way
of a set-associative cache is assigned one Way Guard. The pur-
pose and functionality of these filters are similar to the Segmented
Bloom Filters explained earlier. The only difference is that each
WayGuardkeeps line-fillandreplacement information of thecache
way it is guarding. The following two properties are important in
understanding how the Way Guard technique works:
1. If the filter indicates that the address is not present in the
way it is guarding, then the data is certainly not present in
that particular cache way.
2. If the filter indicates that the address is present, then the data
may be present in the given cache way.
As such, the filter provides a completely safe indication about
the absence of data in the cache way it is guarding. Also, this indi-
cation can be performed within a fixed access time, as opposed to
previously proposed prediction techniques that contain undesirable
variable access times. Figure 3 shows a scenario of an access to
a Way Guard cache. In the example, since the Way Guard filters
of Way 0 and Way 3 indicate a possible hit, only their Tag RAMs
need to be checked, which enables the dynamic energy reduction.
Eventhough thescheme incurs extrahardware, theWayGuardonly
comprises of abit-vector and counters. Querying aWayGuard only
involves checking the bit vector and this consumes much less en-
ergy than looking for the address in the Tag RAM it is guarding.
Note that, since the filter must be checked prior to the Tag RAMs
being selected, there is a potential performance penalty. However,
since the filter is fast, the filter access and the Tag RAM selection
process can be potentially contained within a cycle. In the worst
case, the filter would add one extra cycle to the cache access time.
One implementation variant of the Way Guard technique is illus-
trated by the dotted lines in Figure 3. There are “n” Guards each
guarding a way of an n-way cache. However, note that each filter is
indexed by the same hash function. Given an address, all the filters
will check the same index for the presence of the data. Since for all
cache accesses all the filters are queried, we propose the following
design alternative. In this variant, the vector part of the segmented
Bloom filters are coalesced to form the Way Guard Matrix shown
by the dotted line table. Each row in the matrix contains a bit for
each guard filter. This bit will be the bit in the bit vector corre-
sponding to the index of the row. Thus Matrix[i][k] will consist of
the ithrow of the bit vector of the Way Guard guarding the kth
way. During a cache access, this matrix is first queried as shown
by the dotted arrows coming from the address and the row of bits
obtained for the given index of the address is used to enable only
the cache ways that may contain the data. This technique can fur-
ther reduce energy by using the lower energy matrix structure to
filter out unneeded cache lookups going to the Tag RAMs. The bit
vectors which consist of a very large number of one-bit entries1,
and the majority of the access cost to the bit vector structure goes
in decoding the index. Using one matrix instead of n one bit arrays
saved (n-1) decoder access costs for every access tothe WayGuard.
One notable point about our technique is that it can be extended
toother implementations of highly associative caches such as CAM
tag caches . This can be done by simply adding an AND gate
in the path of every CAM comparator. One input to these AND
gates is obtained from the Way Guard result to effectively reduce
tag comparisons. The detailed evaluation is outside the scope of
We use a modified version of Bochs , a full system x86 em-
ulator, to evaluate the energy savings by taking all effects including
1The number of entries we chose is four times of the number of
cache lines. For a 256KB cache with 32-byte lines, the number of
entries in the array is 32768.
Figure 3: Way Guard Mechanism — Filtering Out Unnecessary Cache Way Lookup
the OS into account. The simulations involved running common
Windows applications like Visual Studio and a set of seven SPEC
benchmarks running on Windows NT on the Bochs emulator.
To collect cache statistics, we integrate a cache simulator en-
hanced with our proposed technique into Bochs, which allows us
to gather cache statistics of various applications directly running on
top of a full OS. Using this framework we can also simulate several
different memory hierarchies simultaneously.
The evaluation for Way Guard is performed in two stages. In
the first stage, only the L2 cache is guarded by the Way Guard fil-
ters. So, we used a fixed size of a 2-way 16KB L1 cache, and 30
different configurations of the L2. Wevaried the capacity by gradu-
ally doubling its size from 64KB up to 2MB. The associativity was
varied in the same manner from 2 to 32 ways.
The size of each Way Guard filter was chosen to be four times
the number of lines of each cache way. We performed experi-
ments with different filter entries and found that having four times
the number of entries gave the best trade-off between filter perfor-
mance and energy savings. The area overhead for all Way Guards
in the L2 is 6% of the L2 size. Notice that this 6% relative over-
head is irrespective of the cache size, as we always use the heuristic
of choosing the Way Guard filter with entries that is four times the
number of entries of the number of cache lines it is guarding.2
We use three bits per counter in the Way Guard counter array.
Since the number of filter entries is four times the number of lines,
it is not expected that for a hash function that can evenly distribute
entries, a three bit counter will overflow. We observed that for all
the experiments performed we did not have a single occurrence of
a counter overflow. We implemented a policy that in the unlikely
event of a counter overflow, we will stop updating entries for that
particular counter and conservatively indicate a possible hit, every
time a query indexes to that particular counter entry.
We show energy savings results for both the serial access and
parallel access versions of the L2 cache. In a serial access, cache
data access follows the tag access only when there is a tag match.
In contrast, for a parallel access cache, the data and tag of all ways
are enabled in parallel, and the correct data is “muxed” out.
2Thereis alsoa small overhead of afew logic gates per WayGuard,
but the overhead is negligible compared to the size of the filter.
In the second stage, only L1 caches are guarded by the Way
Guard. So, we used a 4-way 128KB fixed sized L2 cache, and
20 L1 configurations. Similarly, the L1 capacity was varied from
8KB up to 64KB and their associativity from 2 to 32 ways. The 32-
way L1 is similar to what is employed in XScale processors. We
assume a parallel access L1 cache for all our experiments. We use
the a set of seven SPEC benchmarks that were known to stress the
L2 cache. In addition, we also used five MS Windows applications
used in desktop systems including the booting of Windows NT, Vi-
sual Studio compiling the Bochs source code, an MPEG decoder,
a MP3 decoder, and a simple web browsing application. All the
above applications show sufficient amount of memory activity to
properly illustrate our results. Also, the MS Windows benchmarks
help us understand how the Way Guard technique will behave in a
real multiprocessing environment.
The L1 and L2 caches, the bit vector, and the counter array
were designed using the Artisan 90nm SRAM library in order to
get an estimate for the dynamic and static energy consumption of
the caches and the segmented Bloom filter. The Artisan SRAM
generator is capable of generating synthesizable Verilog code for
SRAMs in 90nm technology. The generated data-sheet provides
the read and write current of the generated SRAM. This gives us
an estimate of the dynamic energy per access of such a structure.
The data-sheet also provides a standby current from which we can
calculate the leakage energy per cycle of the SRAM.
5.3Energy Savings for Way Guard
To illustrate the effectiveness of the Way Guard, we show the av-
erage number of ways looked up (for all the benchmark programs)
for hits and misses for all the 30 L2 cache configurations in Fig-
ure 4 and Figure 5. A notable observation of these results is that
the Way Guard does a very good job in filtering out ways where
the data is not present. In a typical case of a cache hit for an 8-way
cache, only 2.77 ways need to be looked up for a data access. The
average number of ways needed to determine a cache miss is sig-
nificantly lower than that for hits. To determine a miss, the Way
Guard cache checks less than 25% of the ways. Another interest-
ing trend is that the performance of the Way Guards continues to
Direct Mapped 2 ways4 ways 8 ways 16 ways32 ways
Avg # of ways looked up
64 KB128 KB256 KB512 KB1MB2MB
Figure 4: Average Number of Ways Looked Up for Hits in an L2 Cache
Direct Mapped 2 ways4 ways
8 ways16 ways32 ways
Avg # of ways looked up
64 KB 256 KB512 KB1MB 2MB
Figure 5: Average Number of Ways Looked Up for Misses in an L2 Cache
Energy Saving %
2M Cache1M Cache
Figure 6: Energy Savings with respect to Miss Rate
improve with increasing cache sizes. This has to do with the sen-
sitivity of the counting Bloom filter performance with its size. Our
experiments showed that a larger counting Bloom filter always per-
forms better than a smaller one, even though the size of the Bloom
filter is always chosen to be four times the number of cache lines in
each way. For a given associativity, since larger caches have larger
Way Guards guarding their ways, the performance of the filters in
larger caches will be better.
In Figure 6 we try to find out how the total energy savings is af-
fected by the miss rate. The energy savings take into account both
dynamic and leakage energy consumptions of the cache as well
as those consumed by the Way Guards. The baseline is a normal
L2 cache without the Way Guard mechanism. The figure shows
the savings obtained for two cache sizes (1MB and 2MB) for the
“bzip2” from SPEC benchmark. “Bzip2” is chosen for this illus-
tration as a typical benchmark showing trends reflected in all other
benchmarks. We find that the miss rate for a fixed cache size is al-
most the same for associativities greater than 2. Weobserve that we
get significant savings for up to 53% for a 32-way 2 MB cache. As
expected for all cache sizes, the savings increase as the associativ-
ity increases. In other words, the effectiveness of the Way Guard is
increased with higher associativity. The reason behind this is that
the Bloom filters are very effective at indicating absence of data,
and can predict more than 80% of cache misses . In a set asso-
ciative cache lookup, most of the accesses to ways result in misses,
that counting Bloom filters are usually good in predicting. In the
case of a cache hit only one way has the required data and in the
case of a cache miss none of the ways have the data. Thus, a higher
associativity gives a greater chance for indicating absence, leading
to larger energy savings.
By fixing the associativity, the energy savings decrease with in-
creased cachesizes. Onereason for thisis thatfor larger cachesizes
the overhead of the Bloom filter also increases. Also larger caches
contain lower miss rates. So the relative benefits do not completely
account for the larger overheads.
Another trend that was observed was that the relative energy
overhead of Way Guard increases with associativity and decreases
with cache size. We found that for a small 64KB L2 cache the rel-
ative energy overhead of Way Guard ranges from 16% for a 2-way
cache to 22% for a 32-way cache. Instead, for a 2MB cache, the
overhead is only 5% for 2-way to 14% for 32-way. There are two
reasons why the Way Guard’s overhead increases with associativ-
ity. First, as the associativity increases, more bits (one bit for each
way) have to accessed for each access to the cache. Second, as
the associativity increases, the performance of Way Guard also im-
proves leading to less total energy consumption. The reason for the
relative overhead decreasing with increasing cache size is that, as
the cache size increases, the relative energy in accessing the cache
becomes relatively larger than that of accessing the Way Guard.
We compared our technique with the Way Halting technique de-
scribed in . Way Halting uses a fully associative buffer to hold
four tag bits for each line of the cache. When the cache is looked
up, the way halting buffer matches the stored bits for each way of
the corresponding set with the least significant tag bits of the ad-
dress looked up. If these bits do not match for a particular way,
then the lookup will surely miss that way, and the tag comparison
for that way is halted, resulting in energy savings. Weimplemented
the way halting scheme in our Bochs infrastructure. We also mod-
eled the power overheads for the way halting technique using the
Artisan SRAM generator.
The L2 cache energy savings comparing Way Halting and our
Way Guard techniques are shown in Figure 7. The baseline cache
for these relative energy numbers is a serial lookup cache. In a se-
rial lookup cache, to reduce the lookup energy consumption, tag
comparisons and the retrieval of the data portion of a cache line are
done serially, similar to what was described in Figure 3. An access
involves two steps starting with a tag comparison. Only if there is
a tag match will the corresponding data row be accessed to supply
the data line. All benchmark programs show similar energy sav-
ing trends in the serial lookup cache. Thus we report the geometric
64KB 2W 64KB 4W64KB 8W
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1MB 2W 1MB 4W1MB 8W
1MB 16W1MB 32W
2MB 2W2MB 4W 2MB 8W
2MB 16W2MB 32W
Way HaltingWay Guard
Figure 7: Comparing Way Halting and Way Guard Energy Savings in a Serial Lookup Cache
means of energy savings for all the benchmarks in Figure 7. Wesee
that since the data of all the ways are not fired up, the primary sav-
ings with the Way Guard lie in the tag comparisons. We find that
for all cache sizes considered, the Way Guard technique is not very
effective for caches whose associativity is less than 4. The reason
behind this is, for a 2-way cache, Bloom filters save at most one tag
comparison for a hit and 2 tag comparisons for a miss. This bene-
fit in most cases does not surpass the extra energy cost needed for
checking the Bloom filters for each L2 access. In contrast, it shows
energy savings of up to 37% for larger associativity caches. Com-
pared to the Way Halting scheme , the Way Guard technique shows
much better energy savings for 27 of the 30 cache configurations.
In a typical case a 1MB 16-way L2 has a 17% energy savings for a
Way Guard while the Way Halting scheme only achieves 6.3% sav-
ings. Also for low associativities, the Way Halting technique does
not have any energy savings. For a 2MB 4-way cache, Way Halting
technique results in a 7% energy loss. The reason behind this is the
high overhead of Way Halting for every cache access, that involves
comparing four tag bits for every cache way. In contrast, the Way
Guard technique only involves reading “n” bits from a bit vector
array, where “n” is the associativity.
We also compared the energy savings of our technique against
Way Halting based on a parallel lookup cache. The results in Fig-
ure 8 assume that the set-associative cache accesses the same data
row for all cache ways in parallel. It can be easily seen that the
Way Guard technique performs much better than the Way Halting
technique for 25 of the 30 cache configurations. In a typical case,
for a 1 MB 4-way cache, the Way Guard cache shows a saving of
32%, while Way Halting manages to improve the energy by 21%.
As expected, the results for a parallel lookup cache show similar
trends as the results for a serial lookup cache in terms of sensitivity
to cache associativity and cache sizes.
For the L1 cache experiment, we consider the L1 to be a high
performance parallel access cache, where data and tag are accessed
together to achieving a fast hit latency. We applied our Way Guard
technique to both the instruction and data caches. For all our ex-
periments we assume a 2-cycle cache access. We assume that our
Way Guard lookup can be fit into these two cycles to not affect the
L1 cache performance. This assumption is validated by consider-
ing the access times of a typical cache configuration (64KB 4-way)
and its corresponding Way Guard Bloom filter (2048 entries each
way). Using Artisan the access latencies were found to be 0.74ns
for the cache and 0.66ns for the Way Guard. The combined access
time fits into 2 cycles of an embedded processor like AMD Geode
NX 1750 running at 1.4GHz. Note that a normal cache access to
this processor will also take 2 cycles, as the 0.74ns access time to
the cache is larger than the 0.714ns cycle time.
Wefirstshow thegeometric means of thenormalized energy sav-
ings in theL1 I-cache across all thebenchmark inFigure 9. Inthese
experiments, our Way Guard technique shows huge benefits up to
68% for a 32-way cache and more than 50% for a 4-way cache. L1
Figure 9: Average L1 I- and D-Cache Energy Savings
caches show huge benefits because it is almost accessed every cy-
cle during execution. For every access, with the help of Way Guard
filters, only 25 to 30% of the ways need to be checked.
We also performed similar experiments for the L1 D-cache. The
results are also shown in Figure 9. Similar to the I-cache results,
the savings obtained using Way Guard despite the overheads are
very impressive. In a typical case of a 32KB 4-way L1 cache, a
52% overall energy saving was shown.
The initial purpose of Bloom filters was to build memory effi-
cient database applications. Since then, Bloom filters have found
numerous applications in networking and database areas [2, 6, 13,
5]. They were also applied to microarchitectural blocks for predict-
ing load/store collision in LSQ  or optimizing the frequency of
load re-execution .
Theearliestworkof trackingcachemisseswithacounting Bloom
filter is given in . A hardware structure called Jetty was pro-
posed to filter out L2 cache snoops in SMPs. A Jetty-like filter is
also used by Peir et al.  for detecting load misses early in the
pipeline so as to initiate speculative execution. Similarly, Mehta
et al.  also uses a Jetty-like filter to detect L2 misses early to
64KB 2W64KB 4W64KB 8W
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1MB 2W 1MB 4W 1MB 8W
1MB 16W1MB 32W
2MB 2W2MB 4W 2MB 8W
2MB 16W2MB 32W
Way HaltingWay Guard
Figure 8: Comparing Way Halting and Way Guard Energy Savings in a Parallel Lookup Cache
stall the instruction fetch for saving energy. Memik et al.  pro-
posed early cache miss detection hardware encapsulated as Mostly
No Machine (MNM). Their goal is to reduce dynamic cache energy
and toimprove performance by bypassing thecaches that will miss.
Counting Bloom filters were used in  to detect the absence of
synonyms for reducing energy consumption in virtually indexed
caches. We, on the other hand, propose a decoupled Bloom filter
structure where the small bit vector can potentially be kept within
the processor core to perform system dynamic and static energy
conservation of L1 and L2 caches and the core itself.
The most common way prediction mechanism is to predict the
MRU way as proposed in . Similar way prediction techniques
were studied in [9, 12]. Another PC-based way prediction scheme
was proposed in . However, way prediction has the disadvan-
tage of a large performance and energy loss if the prediction is
wrong. One alternative to way prediction is way memoization .
Way memoization keeps way information in the I-cache with valid
bits that ensure that the way information is correct. However, this
technique can only be used in I-caches. Their shortcoming can be
overcome by our Way Guard.
Instead of predicting one single way, way estimation techniques
were proposed to predict a set of ways where the data is guaranteed
to be present for a cache hit. Therefore, way estimation techniques
do not incur a large performance loss for a wrong estimation, be-
cause a wrong estimate only results in a lookup in the cache when
it is missing the cache. One way estimation technique is the sen-
try tag . Way Halting  is an extension of sentry tags. We
did comparisons of this with our Way Guard scheme and showed
much more energy savings in our method than the sentry tags. An-
other way estimationtechnique similar toWayGuard was proposed
in . In which the authors tried to predict lines that have de-
cayed. Since this technique incorporates cache decay, it is not suit-
able for the L1 caches as it may increase the miss rate considerably.
Our technique does not increase the miss rate and thus has no ap-
preciable effect on the performance of the system.
As future applications demand more memory and shrinking fea-
ture sizes allow more one-die transistors, processors are inclined
to incorporate larger caches with higher associativity. These larger
structures, unfortunately, also lead to higher energy consumption.
Thispaper presents anefficientuseof thecounting segmented Bloom
Filtercalled Way Guard to reduce significant dynamic cache energy
by filtering out unnecessary way lookup in a set-associative cache.
Weshowed that our technique can be efficientlyapplied to all levels
of the cache hierarchy, obtaining substantial energy savings of up
to 70% in both instruction and data L1 caches, and up to 65% for a
unified L2 cache. We also showed that our Way Guard outperforms
the recently proposed Way Halting scheme in saving energy.
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