STM Concurrency Control for Embedded Real-TimeSoftware with Tighter Time Bounds

Mohammed El-ShambakeyECE Dept., Virginia Tech

Blacksburg, VA 24060, USA

shambake@vt.edu

Binoy RavindranECE Dept., Virginia Tech

Blacksburg, VA 24060, USA

binoy@vt.edu

ABSTRACT

We consider software transactional memory (STM) concurrencycontrol for multicore real-time software, and present a novel con-tention manager (CM) for resolving transactional conflicts, calledlength-based CM (or LCM). We upper bound transactional retriesand response times under LCM, when used with G-EDF and G-RMA schedulers. We identify the conditions under which LCMoutperforms previous real-time STM CMs and lock-free synchro-nization. Our implementation and experimental studies reveal thatG-EDF/LCM and G-RMA/LCM have shorter or comparable retrycosts and response times than other synchronization techniques.

Categories and Subject Descriptors

C.3 [Special-Purpose and Application-based Systems]: Real-timeand embedded systems

General Terms

Design, Experimentation, Measurement

Keywords

Software transactional memory (STM), real-time contention man-ager

1. INTRODUCTIONLock-based concurrency control suffers from programmability,

scalability, and composability challenges [12]. These challengesare exacerbated in emerging multicore architectures, on which im-proved software performance must be achieved by exposing greaterconcurrency. Transactional memory (TM) is an alternative syn-chronization model for shared memory objects that promises to al-leviate these difficulties. With TM, programmers organize codethat read/write shared objects as transactions, which appear to ex-ecute atomically. Two transactions conflict if they access the sameobject and one access is a write. When that happens, a contentionmanager (or CM) resolves the conflict by aborting one and allowingthe other to commit, yielding (the illusion of) atomicity. In addition

Permission to make digital or hard copies of all or part of this work forpersonal or classroom use is granted without fee provided that copies arenot made or distributed for profit or commercial advantage and that copiesbear this notice and the full citation on the first page. To copy otherwise, torepublish, to post on servers or to redistribute to lists, requires prior specificpermission and/or a fee.DAC 2012, June 3-7, 2012, San Francisco, California, USA.Copyright 2012 ACM 978-1-4503-1199-1/12/06 ...$10.00.

to a simple programming model, TM provides performance com-parable to highly concurrent fine-grained locking and lock-free ap-proaches, and is composable. TM has been proposed in hardware,called HTM, and in software, called STM, with the usual tradeoffs:HTM has lesser overhead, but needs transactional support in hard-ware; STM is available on any hardware. See [11] for an excellentoverview on TM.

Given STM’s programmability, scalability, and composabilityadvantages, we consider it for concurrency control in multicorereal-time software. Doing so requires bounding transactional re-tries, as real-time threads, which subsume transactions, must sat-isfy time constraints. Retry bounds in STM are dependent on theCM policy at hand. Thus, real-time CM is logical.

Past research on real-time CM have proposed resolving trans-actional contention using dynamic and fixed priorities of parentthreads, resulting in Earliest-Deadline-First-based CM (ECM) andRateMonotonic Assignment-based CM (RCM), respectively [7–9].These works show that, ECM and RCM, when used with the GlobalEDF (G-EDF) and Global RMA (G-RMA) multicore schedulers,respectively, achieve higher schedulability than lock-free synchro-nization techniques only under some ranges for the maximum atomicsection length. This raises a fundamental question: is it possible toincrease the atomic section length by an alternative CM design, sothat STM’s schedulability advantage has a larger coverage?

We answer this question by designing a novel CM that can beused with both dynamic and fixed priority (global) multicore real-time schedulers: length-based CM or LCM (Section 4.1). LCMresolves conflicts based on the priority of conflicting jobs, besidesthe length of the interfering atomic section, and the length of theinterfered atomic section. We establish LCM’s retry and responsetime upper bounds, when used with G-EDF (Section 4.2) and withG-RMA (Section 4.5) schedulers. We identify the conditions un-der which G-EDF/LCM outperforms ECM (Section 4.3) and lock-free synchronization (Section 4.4), and G-RMA/LCM outperformsRCM (Section 4.6). We implement LCM and competitor CM tech-niques in the Rochester STM framework [14] and conduct experi-mental studies (Section 5). Our study reveals that G-EDF/LCM andG-RMA/LCM have shorter or comparable retry costs and responsetimes than competitors.

Thus, the paper’s contribution is LCM with superior timelinessproperties. This result thus allows programmers to reap STM’s sig-nificant programmability and composability benefits for a broaderrange of multicore embedded real-time software than what was pre-viously possible.

2. RELATEDWORKTransactional-like concurrency control without using locks, for

real-time systems, has been previously studied in the context of

437

non-blocking data structures (e.g., [1]). Despite their numerous ad-vantages over locks (e.g., deadlock-freedom), their programmabil-ity has remained a challenge. Past studies show that they are bestsuited for simple data structures where their retry cost is competi-tive to the cost of lock-based synchronization [3]. In contrast, STMis semantically simpler [12], and is often the only viable lock-freesolution for complex data structures (e.g., red/black tree) [10] andnested critical sections [15].STM concurrency control for real-time systems has been previ-

ously studied in [2, 7, 9, 10, 13, 16, 17].[13] proposes a restricted version of STM for uniprocessors.

Uniprocessors do not need contention management.[9] bounds response times in distributed systems with STM syn-

chronization. They consider Pfair scheduling, limit to small atomicregions with fixed size, and limit transaction execution to span atmost two quanta. In contrast, we allow transaction lengths witharbitrary duration.

[16] presents real-time scheduling of transactions and serializestransactions based on deadlines. However, the work does not boundretries and response times. In contrast, we establish such bounds.

[17] proposes real-time HTM. The work does not describe howtransactional conflicts are resolved. Besides, the retry bound as-sumes that the worst case conflict between atomic sections of dif-ferent tasks occurs when the sections are released at the same time.However, we show that this is not the worst case. We develop retryand response time upper bounds based on much worse conditions.

[10] upper bounds retries and response times for ECM with G-EDF, and identify the tradeoffs with locking and lock-free proto-cols. Similar to [17], [10] also assumes that the worst case conflictbetween atomic sections occurs when the sections are released si-multaneously. The ideas in [10] are extended in [2], which presentsthree real-time CM designs. But no retry bounds or schedulabilityanalysis techniques are presented for those CMs.

[7] presents the ECM and RCM contention managers, and upperbounds transactional retries and response times under them. Thework also identifies the conditions under which ECM and RCM aresuperior to lock-free synchronization. In particular, they show that,STM’s superiority holds only under some ranges for the maximumatomic section length. Our work builds upon this result.

3. PRELIMINARIESWe consider a multiprocessor system withm identical processors

and n sporadic tasks τ1,τ2, . . . ,τn. The kth instance (or job) of atask τi is denoted τ

ki . Each task τi is specified by its worst case

execution time (WCET) ci, its minimum period Ti between any twoconsecutive instances, and its relative deadline Di, where Di = Ti.

Job τji is released at time r

ji and must finish no later than its absolute

deadline dji = r

ji +Di. Under a fixed priority scheduler such as G-

RMA, pi determines τi’s (fixed) priority and it is constant for allinstances of τi. Under a dynamic priority scheduler such as G-

EDF, a job τji ’s priority, p

ji , differs from one instance to another. A

task τ j may interfere with task τi for a number of times during aninterval L, and this number is denoted as Gi j(L).Shared objects. A task may need to read/write shared, in-memory

data objects while it is executing any of its atomic sections, whichare synchronized using STM. The set of atomic sections of task τi

is denoted si. ski is the k

th atomic section of τi. Each object, θ, canbe accessed by multiple tasks. The set of distinct objects accessedby τi is θi without repeating objects. The set of atomic sectionsused by τi to access θ is si(θ), and the sum of the lengths of thoseatomic sections is len(si(θ)). s

ki (θ) is the kth atomic section of τi

that accesses θ. ski (θ) executes for a duration len(ski (θ)). The set of

tasks sharing θ with τi is denoted γi(θ).Atomic sections are non-nested (supporting nested STM is future

work). Each section is assumed to access only one object; thisallows us to be consistent with the assumptions in [7], enablinga comparison with retry-loop lock-free synchronization [5], whichis an important goal of this paper. The maximum-length atomicsection in τi that accesses θ is denoted simax(θ), while the maximumone among all tasks is smax(θ), and the maximum one among taskswith priorities lower than that of τi is s

imax(θ).

STM retry cost. If two or more atomic sections conflict, the CMwill commit one section and abort and retry the others, increasingthe time to execute the aborted sections. The increased time that anatomic section s

pi (θ) will take to execute due to a conflict with an-

other section skj(θ), is denotedWpi (s

kj(θ)). If an atomic section, s

pi ,

is already executing, and another atomic section skj tries to access

a shared object with spi , then skj is said to “interfere" or “conflict"

with spi . The job skj is the “interfering job", and the job s

pi is the

“interfered job." The total time that a task τi’s atomic sections haveto retry over Ti is denoted RC(Ti). The additional amount of timethat a task τ j causes to response time of τi when interfering withτi during L, without considering retries due to atomic sections, isdenotedWi j(L).

4. LENGTH-BASED CMLCM resolves conflicts based on the priority of conflicting jobs,

besides the length of the interfering atomic section, and the lengthof the interfered atomic section. This is in contrast to ECM andRCM [7], where conflicts are resolved using the priority of the con-flicting jobs. This strategy allows lower priority jobs, under LCM,to retry for lesser time than that under ECM and RCM, but higherpriority jobs, sometimes, wait for lower priority ones with boundedpriority-inversion.

4.1 Design and Rationale

Algorithm 1: LCM

Data: ski (θ)→ interfered atomic section.

slj(θ)→ interfering atomic section.

ψ → predefined threshold ∈ [0,1].

δki (θ)→ remaining execution length of ski (θ)

Result: which atomic section of ski (θ) or slj(θ) aborts

1 if pki > plj then

2 slj(θ) aborts;

3 else

4 ckli j = len(slj(θ))/len(ski (θ));

5 αkli j = ln(ψ)/(ln(ψ)− ckli j );

6 α =(

len(ski (θ))−δki (θ)

)

/len(ski (θ));

7 if α ≤ αkli j then

8 ski (θ) aborts;9 else

10 slj(θ) aborts;

11 end12 end

For both ECM and RCM, ski (θ) can be totally repeated if slj(θ)

— which belongs to a higher priority job τbj than τ

ai — conflicts

with ski (θ) at the end of its execution, while ski (θ) is just about tocommit. Thus, LCM, shown in Algorithm 1, uses the remaining

438

length of ski (θ)when it is interfered, as well as len(slj(θ)), to decide

which transaction must be aborted. If pki was greater than plj, then

ski (θ) would be the one that commits, because it belongs to a higher

priority job, and it started before slj(θ) (step 2). Otherwise, ckli j is

calculated (step 4) to determine whether it is worth aborting ski (θ)

in favor of slj(θ), because len(slj(θ)) is relatively small compared

to the remaining execution length of ski (θ) (explained further).We assume that:

ckli j = len(slj(θ))/len(ski (θ)) (1)

where ckli j ∈]0,∞[, to cover all possible lengths of slj(θ). Our idea

is to reduce the opportunity for the abort of ski (θ) if it is close to

committing when interfered and len(slj(θ)) is large. This abort op-

portunity is increasingly reduced as ski (θ) gets closer to the end of

its execution, or len(slj(θ)) gets larger.

On the other hand, as ski (θ) is interfered early, or len(slj(θ)) is

small compared to ski (θ)’s remaining length, the abort opportunity

is increased even if ski (θ) is close to the end of its execution. To

decide whether ski (θ) must be aborted or not, we use a threshold

value ψ ∈ [0,1] that determines αkli j (step 5), where α

kli j is the max-

imum percentage of len(ski (θ)) below which slj(θ) is allowed to

abort ski (θ). Thus, if the already executed part of ski (θ) — when

slj(θ) interferes with ski (θ) — does not exceed αkli j len(s

ki (θ)), then

ski (θ) is aborted (step 8). Otherwise, slj(θ) is aborted (step 10).

0

0. 2

0. 4

0. 6

0. 8

1

0 0. 2 0. 4 0. 6 0. 8 1 1. 2 1. 4

Abort percent

s i per cent ( α)

sj=0.01sisj=0.1si

sj=sisj=10si

sj=100si

α1 α2 α3

�

Figure 1: Interference of ski (θ) by various lengths of slj(θ)

The behavior of LCM is illustrated in Figure 1. In this figure, thehorizontal axis corresponds to different values of α ranging from 0to 1, and the vertical axis corresponds to different values of abortopportunities, f (ckli j ,α), ranging from 0 to 1 and calculated by (2):

f (ckli j ,α) = e−ckl

i jα

1−α (2)

where ckli j is calculated by (1).

Figure 1 shows one atomic section ski (θ) (whose α changes along

the horizontal axis) interfered by five different lengths of slj(θ). For

a predefined value of f (ckli j ,α) (denoted as ψ in Algorithm 1), there

corresponds a specific value of α (which is αkli j in Algorithm 1)

for each curve. For example, when len(slj(θ)) = 0.1× len(ski (θ)),

slj(θ) aborts ski (θ) if the latter has not executed more than α3 per-

centage (shown in Figure 1) of its execution length. As len(slj(θ))

decreases, the corresponding αkli j increases (as shown in Figure 1,

α3> α2> α1).Equation (2) achieves the desired requirement that the abort op-

portunity is reduced as ski (θ) gets closer to the end of its execution

(as α → 1, f (ckli j ,1)→ 0), or as the length of the conflicting trans-

action increases (as ckli j → ∞, f (∞,α)→ 0). Meanwhile, this abort

opportunity is increased as ski (θ) is interfered closer to its release

(as α → 0, f (ckli j ,0)→ 1), or as the length of the conflicting trans-

action decreases (as ckli j → 0, f (0,α)→ 1).LCM is not a centralized CM, which means that, upon a conflict,

each transactions has to decide whether it must commit or abort.

CLAIM 1. Let slj(θ) interfere once with ski (θ) at αkli j . Then, the

maximum contribution of slj(θ) to ski (θ)’s retry cost is:

W ki (s

lj(θ))≤ α

kli j len

(

ski (θ))

+ len(

slj(θ))

(3)

(Proofs of all claims are provided in the Supplementary Materialsection at the end of the paper.)

CLAIM 2. An atomic section of a higher priority job, τbj , may

have to abort and retry due to a lower priority job, τai , if s

lj(θ)

interferes with ski (θ) after the αkli j percentage. τ j’s retry time, due

to ski (θ) and slj(θ), is upper bounded by:

W lj (s

ki (θ))≤

(

1−αkli j

)

len(

ski (θ))

(4)

CLAIM 3. A higher priority job, τzi , suffers from priority inver-

sion for at most number of atomic sections in τzi .

CLAIM 4. The maximum delay suffered by slj(θ) due to priorityinversion is caused by the maximum length atomic section access-

ing object θ, which belongs to a lower priority job than τbj that

owns slj(θ).

4.2 Response Time of G-EDF/LCM

CLAIM 5. RC(Ti) for a task τi under G-EDF/LCM is upper

bounded by:

RC(Ti) =

(

∑∀τh∈γi

∑∀θ∈θi∧θh

(

⌈

Ti

Th

⌉

∑∀slh(θ)

len(

slh(θ))

+ αhlmaxlen

(

shmax(θ))

))

+ ∑∀syi (θ)

(

1−αiymax

)

len(

simax(θ))

(5)

where αhlmax is the α value that corresponds to ψ due to the inter-

ference of shmax(θ) by slh(θ). α

iymax is the α value that corresponds

to ψ due to the interference of simax(θ) by syi (θ).

Response time of τi is calculated by (11) in [7].

439

Jh

h

i

2

h

3

h T T

t(Ti)

1

h

p

hp+1

h

τ

ττ τ τ τ τ

Ti

Figure 2: τph has a higher priority than τ

xi

4.3 Schedulability of G-EDF/LCM and ECMWe now compare the schedulability of G-EDF/LCMwith ECM [7]

to understand when G-EDF/LCM will perform better. Toward this,we compare the total utilization of ECM with that of G-EDF/LCM.For each method, we inflate the ci of each task τi by adding the retrycost suffered by τi. Thus, if method A adds retry cost RCA(Ti) to ci,and method B adds retry cost RCB(Ti) to ci, then the schedulabilityof A and B are compared as:

∑∀τi

ci+RCA(Ti)

Ti≤ ∑

∀τi

ci+RCB(Ti)

Ti

∑∀τi

RCA(Ti)

Ti≤ ∑

∀τi

RCB(Ti)

Ti(6)

Thus, schedulability is compared by substituting the retry cost addedby the synchronization methods in (6).

CLAIM 6. Let smax be the maximum length atomic section ac-

cessing any object θ. Let αmax and αmin be the maximum and

minimum values of α for any two atomic sections ski (θ) and slj(θ).Given a threshold ψ, schedulability of G-EDF/LCM is equal or bet-

ter than ECM if for any task τi:

1−αmin

1−αmax≤ ∑

∀τh∈γi

⌈

Ti

Th

⌉

(7)

4.4 G-EDF/LCM versus Lock-freeWe consider the retry-loop lock-free synchronization for G-EDF

given in [5]. This lock-free approach is the most relevant to ourwork.

CLAIM 7. Let smax denote len(smax) and rmax denote the max-

imum execution cost of a single iteration of any retry loop of any

task in the retry-loop lock-free algorithm in [5]. Now, G-EDF/LCM

achieves higher schedulability than the retry-loop lock-free approach

if the upper bound on smax/rmax under G-EDF/LCM ranges be-

tween 0.5 and 2 (which is higher than that under ECM).

4.5 Response Time of G-RMA/LCM

CLAIM 8. Let

λ2( j,θ) = ∑∀slj(θ)

len(slj(θ))+αjlmaxlen(s

jmax(θ))

where αjlmax is the α value corresponding to ψ due to the inter-

ference of sjmax(θ) by slj(θ). The retry cost of any task τi under

G-RMA/LCM during Ti is given by:

RC (Ti) = ∑∀τ

∗j

∑θ∈(θi∧θ j)

((⌈

Ti

Tj

⌉

+1

)

λ2( j,θ)

)

+ ∑∀syi (θ)

(

1−αiymax

)

len(

simax(θ))

(8)

where τ∗j = {τ j|(τ j ∈ γi)∧ (p j > pi)}.

The response time is calculated by (17) in [7] with replacingRC(R

upi ) with RC(Ti).

4.6 Schedulability of G-RMA/LCMandRCM

CLAIM 9. Under the same assumptions of Claims 6 and 8, G-

RMA/LCM’s schedulability is equal or better than RCM if:

1−αmin

1−αmax≤ ∑

∀τ∗j

(⌈

Ti

Tj

⌉

+1

)

(9)

5. EXPERIMENTAL EVALUATIONHaving established LCM’s retry and response time upper bounds,

and the conditions under which it outperforms ECM, RCM, andlock-free synchronization, we now would like to understand howLCM’s retry and response times in practice (i.e., on average) com-pare with that of competitor methods. Since this can only be un-derstood experimentally, we implement LCM and the competitormethods and conduct experimental studies.

5.1 Experimental SetupWe used the ChronOS real-time Linux kernel [4] and the RSTM

library [14]. We modified RSTM to include implementations ofECM, RCM, G-EDF/LCM, and G-RMA/LCM contention managers,and modified ChronOS to include implementations of G-EDF andG-RMA schedulers.

For the retry-loop lock-free implementation, we used a loop thatreads an object and attempts to write to the object using a compare-and-swap (CAS) instruction. The task retries until the CAS suc-ceeds.

We use an 8 core, 2GHz AMD Opteron platform. The aver-age time taken for one write operation by RSTM on any core is0.0129653375µs, and the average time taken by one CAS-loop op-eration on any core is 0.0292546250 µs.

We used the periodic task set shown in Table 1. Each task runsin its own thread and has an atomic section. Atomic section prop-erties are probabilistically controlled (for experimental evaluation)using three parameters: the maximum and minimum lengths of anyatomic section within the task, and the total length of atomic sec-tions within any task. All task atomic sections access the sameobject, and do write operations on the object (thus, contention isthe highest).

5.2 ResultsFigure 3 shows the retry cost (RC) for each task in the three

task sets in Table 1, where each task’s atomic section length isequal to half of the task length. Each data point in the figure hasa confidence level of 0.95. We observe that G-EDF/LCM and G-RMA/LCM achieve shorter or comparable retry cost than ECM andRCM. Since all tasks are initially released at the same time, and dueto the specific nature of task properties, tasks with lower IDs some-how have higher priorities under the G-EDF scheduler. Note thattasks with lower IDs have higher priorities under G-RMA, sincetasks are ordered in non-decreasing order of their periods.

Thus, we observe that G-EDF/LCM and G-RMA/LCM achievecomparable retry costs to ECM and RCM for some tasks with lowerIDs. But when task ID increases, LCM — for both schedulers —achieves much shorter retry costs than ECM and RCM. This is be-cause, higher priority tasks in LCM suffers blocking by lower pri-ority tasks, which is not the case for ECM and RCM. However, astask priority decreases, LCM, by definition, prevents higher prior-ity tasks from aborting lower priority ones if a higher priority task

440

Table 1: Task sets. (a) Task set 1: 5-task set; (b) Task set 2:

10-task set; (c) Task set 3: 12-task set

(a)Ti(µs) ci(µs)

τ1 500000 150000

τ2 1000000 227000

τ3 1500000 410000

τ4 3000000 299000

τ5 5000000 500000

(b)Ti(µs) ci(µs)

τ1 400000 75241

τ2 750000 69762

τ3 1200000 267122

τ4 1500000 69863

τ5 2400000 152014

τ6 4000000 286301

τ7 7500000 493150

τ8 10000000 794520

τ9 15000000 1212328

τ10 20000000 1775342

(c)Ti(µs) ci(µs)

τ1 400000 58195

τ2 750000 53963

τ3 1000000 206330

τ4 1200000 53968

τ5 1500000 117449

τ6 2400000 221143

τ7 3000000 290428

τ8 4000000 83420

τ9 7500000 380917

τ10 10000000 613700

τ11 15000000 936422

τ12 20000000 1371302

interferes with a lower priority one after a specified threshold. Incontrast, under ECM and RCM, lower priority tasks abort in fa-vor of higher priority ones. G-EDF/LCM and G-RMA/LCM alsoachieve shorter retry costs than the retry-loop lock-free algorithm.Figure 4 shows the response time of each task in the Table 1 task

sets with a confidence level of 0.95. (Again, each task’s atomic sec-tion length is equal to half of the task length.) We observe that G-EDF/LCM and G-RMA/LCM achieve shorter response time thanthe retry-loop lock-free algorithm, and shorter or comparable re-sponse time than ECM and RCM.We repeated the experiments by varying three parameters: the

relative total length of all atomic sections to the length of the task,the maximum relative length of any atomic section to the lengthof the task, and the minimum relative length of any atomic sectionto the length of the task. Full set of results are omitted here dueto space constraints; however additional results are included in theSupplementary Material section. Full set of results are given inAppendix B in [6].

6. CONCLUSIONSIn ECM and RCM, a task incurs at most 2smax retry cost for

each of its atomic section due to conflict with another task’s atomicsection. With LCM, this retry cost is reduced to (1+αmax)smaxfor each aborted atomic section. In ECM and RCM, tasks do notretry due to lower priority tasks, whereas in LCM, they do so. InG-EDF/LCM, retry due to a lower priority job is encountered onlyfrom a task τ j’s last job instance during τi’s period. This is notthe case with G-RMA/LCM, because, each higher priority taskcan be aborted and retried by any job instance of lower prioritytasks. Schedulability of G-EDF/LCM and G-RMA/LCM is bet-ter or equal to ECM and RCM, respectively, by proper choices for

-2e+08

0

2e+08

4e+08

6e+08

8e+08

1e+09

1.2e+09

1.4e+09

1.6e+09

1.8e+09

2e+09

0.5 1 1.5 2 2.5 3 3.5 4 4.5 5 5.5

RC

(nsec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(a) Task set 1

0

1e+09

2e+09

3e+09

4e+09

5e+09

6e+09

7e+09

8e+09

9e+09

1e+10

0 1 2 3 4 5 6 7 8 9 10 11

RC

(nsec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(b) Task set 2

-2e+16

0

2e+16

4e+16

6e+16

8e+16

1e+17

1.2e+17

0 2 4 6 8 10 12 14

RC

(nsec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(c) Task set 3

Figure 3: Task retry costs under LCM and competitor synchro-

nization methods

441

0

5e+08

1e+09

1.5e+09

2e+09

2.5e+09

3e+09

3.5e+09

4e+09

4.5e+09

0.5 1 1.5 2 2.5 3 3.5 4 4.5 5 5.5

Res-T

ime(n

sec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(a) Task set 1

0

2e+09

4e+09

6e+09

8e+09

1e+10

1.2e+10

1.4e+10

0 1 2 3 4 5 6 7 8 9 10 11

Res-T

ime(n

sec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(b) Task set 2

0

1e+09

2e+09

3e+09

4e+09

5e+09

6e+09

7e+09

8e+09

9e+09

0 2 4 6 8 10 12 14

Res-T

ime(n

sec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(c) Task set 3

Figure 4: Task response times under LCM and competitor syn-

chronization methods

αmin and αmax. Schedulability of G-EDF/LCM is better than retry-loop lock-free synchronization for G-EDF if the upper bound onsmax/rmax is between 0.5 and 2, which is higher than that achievedby ECM.

Acknowledgments

This work is supported in part by NSF CNS 0915895, NSF CNS1116190, and NSF CNS 1130180.

7. REFERENCES[1] J. Anderson, S. Ramamurthy, and K. Jeffay. Real-time computing

with lock-free shared objects. In RTSS, pages 28–37, 1995.

[2] A. Barros and L. Pinho. Managing contention of softwaretransactional memory in real-time systems. In IEEE RTSS,

Work-In-Progress, 2011.

[3] B. B. Brandenburg et al. Real-time synchronization onmultiprocessors: To block or not to block, to suspend or spin? InRTAS, pages 342–353, 2008.

[4] M. Dellinger, P. Garyali, and B. Ravindran. Chronos linux: abest-effort real-time multiprocessor linux kernel. In Design

Automation Conference (DAC), 2011 48th ACM/EDAC/IEEE, pages474–479. IEEE, 2011.

[5] U. C. Devi, H. Leontyev, and J. H. Anderson. Efficientsynchronization under global EDF scheduling on multiprocessors. InECRTS, pages 75–84, 2006.

[6] M. El-Shambakey and B. Ravindran. On the design of real-time stmcontention managers. Technical report, ECE Department, VirginiaTech, 2011. Available as http://www.real-time.ece.vt.edu/tech-report-rt-stm-cm11.pdf.

[7] M. Elshambakey and B. Ravindran. Stm concurrency control formulticore embedded real-time software: Time bounds and tradeoffs.In SAC, 2012.

[8] S. Fahmy, B. Ravindran, and E. Jensen. Response time analysis ofsoftware transactional memory-based distributed real-time systems.In ACM SAC, pages 334–338, 2009.

[9] S. Fahmy, B. Ravindran, and E. D. Jensen. On bounding responsetimes under software transactional memory in distributedmultiprocessor real-time systems. In DATE, pages 688–693, 2009.

[10] S. F. Fahmy. Collaborative Scheduling and Synchronization ofDistributable Real-Time Threads. PhD thesis, Virginia Tech, 2010.

[11] T. Harris, J. Larus, and R. Rajwar. Transactional Memory. Morgan &Claypool Publishers, 2nd. edition, December 2010.

[12] M. Herlihy. The art of multiprocessor programming. In PODC, pages1–2, 2006.

[13] J. Manson, J. Baker, et al. Preemptible atomic regions for real-timeJava. In RTSS, pages 10–71, 2006.

[14] V. Marathe, M. Spear, C. Heriot, A. Acharya, D. Eisenstat,W. Scherer III, and M. Scott. Lowering the overhead of nonblockingsoftware transactional memory. In Workshop on Languages,

Compilers, and Hardware Support for Transactional Computing

(TRANSACT), 2006.

[15] B. Saha, A.-R. Adl-Tabatabai, et al. McRT-STM: a high performancesoftware transactional memory system for a multi-core runtime. InPPoPP, pages 187–197, 2006.

[16] T. Sarni, A. Queudet, and P. Valduriez. Real-time support forsoftware transactional memory. In RTCSA, pages 477–485, 2009.

[17] M. Schoeberl, F. Brandner, and J. Vitek. RTTM: Real-timetransactional memory. In ACM SAC, pages 326–333, 2010.

Supplementary Material

This section includes proofs of all Claims.

S.1 Proof of Claim 1

PROOF. If slj(θ) interferes with ski (θ) at a ϒ percentage, where

ϒ < αkli j , then the retry cost of ski (θ) is ϒlen(ski (θ)) + len(slj(θ)),

which is lower than that calculated in (3). Besides, if slj(θ) inter-

feres with ski (θ) after αkli j percentage, then s

ki (θ) will not abort.

442

S.2 Proof of Claim 2

PROOF. It is derived directly from Claim 1, as slj(θ) will have

to retry for the remaining length of ski (θ).

S.3 Proof of Claim 3

PROOF. Assuming three atomic sections, ski (θ), slj(θ) and s

ba(θ),

where p j > pi and slj(θ) interferes with s

ki (θ) after α

kli j . Then s

lj(θ)

will have to abort and retry. At this time, if sba(θ) interferes with theother two atomic sections, and the LCM decides which transactionto commit based on comparison between each two transactions. So,we have the following cases:-

• pa < pi < p j , then sba(θ) will not abort any one because it isstill in its beginning and it is of the lowest priority. So. τ j isnot indirectly blocked by τa.

• pi < pa < p j and even if sba(θ) interferes with ski (θ) before

αkbia , so, s

ba(θ) is allowed abort ski (θ). Comparison between

slj(θ) and sba(θ) will result in LCM choosing slj(θ) to commit

and abort sba(θ) because the latter is still beginning, and τ j is

of higher priority. If sba(θ) is not allowed to abort ski (θ), the

situation is still the same, because slj(θ) was already retrying

until ski (θ) finishes.

• pa > p j > pi, then if sba(θ) is chosen to commit, this is notpriority inversion for τ j because τa is of higher priority.

• if τa preempts τi, then LCM will compare only between slj(θ)

and sba(θ). If pa < p j, then slj(θ) will commit because of its

task’s higher priority and sba(θ) is still at its beginning, oth-erwise, slj(θ) will retry, but this will not be priority inversionbecause τa is already of higher priority than τ j . If τa doesnot access any object but it preempts τi, then CM will chooseslj(θ) to commit as only already running transactions are com-peting together.

So, by generalizing these cases to any number of conflicting jobs, itis seen that when an atomic section, slj(θ), of a higher priority jobis in conflict with a number of atomic sections belonging to lowerpriority jobs, slj(θ) can suffer from priority inversion by only oneof them. So, each higher priority job can suffer priority inversionat most its number of atomic section. Claim follows.

S.4 Proof of Claim 4

PROOF. Assume three atomic sections, ski (θ), slj(θ), and szh(θ),

where p j > pi, p j > ph, and len(ski (θ)) > len(szh(θ)). Now, αkli j >

αzlh j and ckli j < czlh j . By applying (4) to obtain the contribution of

ski (θ) and szh(θ) to the priority inversion of s

lj(θ) and dividing them,

we get:

W lj (s

ki (θ))

W lj (s

zh(θ))

=

(

1−αkli j

)

len(ski (θ))(

1−αzlh j

)

len(szh(θ))

By substitution for αs from (2):

=(1− lnψ

lnψ−ckli j)len(ski (θ))

(1− lnψ

lnψ−czlh j)len(szh(θ))

=(

−ckli j

lnψ−ckli j)len(ski (θ))

(−czlh j

lnψ−czlh j)len(szh(θ))

∵ lnψ ≤ 0 and ckli j ,cklh j > 0,∴ by substitution from (1)

=len(slj(θ))/(lnψ− ckli j )

len(slj(θ))/(lnψ− czlh j)=

lnψ− czlh j

lnψ− ckli j> 1

Thus, as the length of the interfered atomic section increases, the ef-fect of priority inversion on the interfering atomic section increases.Claim follows.

S.5 Proof of Claim 5

PROOF. The maximum number of higher priority instances of

τh that can interfere with τxi is

⌈

TiTh

⌉

, as shown in Figure 2, where

one instance of τh and τph coincides with the absolute deadline of

τxi .By using Claims 1, 2, 3, and 4, and Claim 1 in [7] to determine

the effect of atomic sections belonging to higher and lower priorityinstances of interfering tasks to τ

xi , claim follows.

S.6 Proof of Claim 6

PROOF. Under ECM, RC(Ti) is upper bounded by:

RC(Ti)≤ ∑∀τh∈γi

∑∀θ∈ (θi∧θh)

⌈

Ti

Th

⌉

∑∀szh(θ)

2len(smax)

(10)

with the assumption that all lengths of atomic sections of (4) and(8) in [7] and (5) are replaced by smax. Let α

hlmax in (5) be replaced

with αmax, and αiymax in (5) be replaced with αmin. As αmax, αmin,

and len(smax) are all constants, (5) is upper bounded by:

RC(Ti) ≤

(

∑∀τh∈γi

∑∀θ∈θi∧θh

(

⌈

Ti

Th

⌉

∑∀slh(θ)

(1+αmax)

len(

smax

)

))

+ ∑∀syi (θ)

(

1−αmin

)

len(

smax

)

(11)

If βih1 is the total number of times any instance of τh accesses shared

objects with τi, then βih1 = ∑∀θ∈(θi∧θh) ∑∀szh(θ)

. Furthermore, if βi2

is the total number of times any instance of τi accesses shared ob-jects with any other instance, β

i2 = ∑∀syi (θ)

, where θ is shared with

another task. Then, βi =max{max∀τh∈γi{β

ih1 },β

i2} is the maximum

number of accesses to all shared objects by any instance of τi or τh.Thus, (10) becomes:

RC(Ti)≤ ∑τh∈γi

2

⌈

Ti

Th

⌉

βilen(smax) (12)

and (11) becomes:

RC(Ti) ≤ βilen(smax)

(

(1−αmin)

+ ∑∀τh∈γi

⌈

Ti

Th

⌉

(1+αmax)

)

(13)

We can now compare the total utilization of G-EDF/LCM with

443

that of ECM by comparing (11) and (13) for all τi:

∑∀τi

(1−αmin)+∑∀τh∈γi

(⌈

TiTh

⌉

(1+αmax))

Ti

≤ ∑∀τi

∑∀τh∈γi2⌈

TiTh

⌉

Ti(14)

(14) is satisfied if for each τi, the following condition is satisfied:

(1−αmin)+ ∑∀τh∈γi

(⌈

Ti

Th

⌉

(1+αmax)

)

≤ 2 ∑∀τh∈γi

⌈

Ti

Th

⌉

∴ 1−αmin

1−αmax≤ ∑

∀τh∈γi

⌈

Ti

Th

⌉

Claim follows.

S.7 Proof of Claim 7

PROOF. From [5], the retry-loop lock-free algorithm is upperbounded by:

RL(Ti) = ∑τh∈γi

(⌈

Ti

Th

⌉

+1

)

βirmax (15)

where βi is as defined in Claim 6. The retry cost of τi in G-EDF/LCM is upper bounded by (13). By comparing G-EDF/LCM’stotal utilization with that of the retry-loop lock-free algorithm, weget:

∑∀τi

(

(1−αmin)+∑∀τh∈γi

(⌈

TiTh

⌉

(1+αmax)))

βismax

Ti

≤ ∑∀τi

∑∀τh∈γi

(⌈

TiTh

⌉

+1)

βirmax

Ti

∴ smax

rmax≤

∑∀τi

∑∀τh∈γi

(⌈

TiTh

⌉

+1)

βi

Ti

∑∀τi

(

(1−αmin)+∑∀τh∈γi

(⌈

TiTh

⌉

(1+αmax)))

βi

Ti

(16)

Let the number of tasks that have shared objects with τi be ω (i.e.,

∑τh∈γi= ω ≥ 1 since at least one task has a shared object with τi;

otherwise, there is no conflict between tasks). Let the total number

of tasks be n, so 1 ≤ ω ≤ n− 1, and⌈

TiTh

⌉

∈ [1,∞[. To find the

minimum and maximum values for the upper bound on smax/rmax,we consider the following cases:

• αmin → 0,αmax → 0

∴ (16) will be:

smax

rmax≤ 1+

∑∀τi

ω−1Ti

∑∀τi

1+∑∀τh∈γi

⌈

TiTh

⌉

Ti

(17)

By substituting the edge values for ω and⌈

TiTh

⌉

in (17), we derive

that the upper bound on smax/rmax lies between 1 and 2.

• αmin → 0,αmax → 1

(16) becomes

smax

rmax≤ 0.5+

∑∀τi

ω−0.5Ti

∑∀τi

1+2∑∀τh∈γi

⌈

TiTh

⌉

Ti

(18)

By applying the edge values for ω and⌈

TiTh

⌉

in (18), we derive that

the upper bound on smax/rmax lies between 0.5 and 1.

• αmin → 1,αmax → 0

This case is rejected since αmin ≤ αmax.

• αmin → 1,αmax → 1

∴ (16) becomes:

smax

rmax≤ 0.5+

∑τi

ω

Ti

2∑τi

∑∀τh∈γi

⌈

TiTh

⌉

Ti

(19)

By applying the edge values for ω and⌈

TiTh

⌉

in (19), we derive

that the upper bound on smax/rmax lies between 0.5 and 1, which issimilar to that achieved by ECM.

Summarizing from the previous cases, the upper bound on smax/rmaxlies between 0.5 and 2, whereas for ECM [7], it lies between 0.5and 1. Claim follows.

S.8 Proof of Claim 8

PROOF. Under G-RMA, all instances of a higher priority task,τ j , can conflict with a lower priority task, τi, during Ti. (3) canbe used to determine the contribution of each conflicting atomicsection in τ j to τi. Meanwhile, all instances of any task with lowerpriority than τi can conflict with τi during Ti. Claims 2 and 3 can beused to determine the contribution of conflicting atomic sections inlower priority tasks to τi. Using the previous notations and Claim3 in [7], the claim follows.

S.9 Proof of Claim 9

PROOF. Under the same assumptions as that of Claims 6 and 8,(8) can be upper bounded as:

RC(Ti) ≤ ∑∀τ

∗j

((⌈

Ti

Tj

⌉

+1

)

(1+αmax)len(smax)βi

)

+ (1−αmin)len(smax)βi (20)

For RCM, (16) in [7] for RC(Ti) is upper bounded by:

RC(Ti)≤ ∑∀τ

∗j

(⌈

Ti

Tj

⌉

+1

)

2βilen(smax)

By comparing the total utilization of G-RMA/LCM with that ofRCM, we get:

∑∀τi

len(smax)βi

(

(1−αmin)+∑∀τ∗j

((⌈

TiTj

⌉

+1)

(1+αmax)))

Ti

≤ ∑∀τi

2len(smax)βi ∑∀τ∗j

(⌈

TiTj

⌉

+1)

Ti(21)

(21) is satisfied if ∀τi (9) is satisfied. Claim follows.

S.10 EXTENDED RESULTSThe three parameters x,y,z for each figure specify respectively

the relative total length of all atomic sections to the length of thetask, the maximum relative length of any atomic section to thelength of the task, and the minimum relative length of any atomicsection to the length of the task.

444

0

2e+08

4e+08

6e+08

8e+08

1e+09

1.2e+09

1.4e+09

1.6e+09

1.8e+09

2e+09

0.5 1 1.5 2 2.5 3 3.5 4 4.5 5 5.5

RC

(nsec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(a) Task set 1

0

1e+09

2e+09

3e+09

4e+09

5e+09

6e+09

7e+09

8e+09

9e+09

0 1 2 3 4 5 6 7 8 9 10 11

RC

(nse

c)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(b) Task set 2

0

2e+16

4e+16

6e+16

8e+16

1e+17

1.2e+17

0 2 4 6 8 10 12 14

RC

(nse

c)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(c) Task set 3

Figure 5: Task retry costs under LCM and competitor synchro-

nization methods (0.5,0.2,0.2)

0

5e+08

1e+09

1.5e+09

2e+09

2.5e+09

3e+09

3.5e+09

4e+09

0.5 1 1.5 2 2.5 3 3.5 4 4.5 5 5.5

Res-T

ime(n

sec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(a) Task set 1

0

2e+09

4e+09

6e+09

8e+09

1e+10

1.2e+10

0 1 2 3 4 5 6 7 8 9 10 11

Res-T

ime(n

sec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(b) Task set 2

0

1e+09

2e+09

3e+09

4e+09

5e+09

6e+09

7e+09

8e+09

9e+09

0 2 4 6 8 10 12 14

Res-T

ime(n

sec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(c) Task set 3

Figure 6: Task response times under LCM and competitor syn-

chronization methods (0.5,0.2,0.2)

445

0

5e+08

1e+09

1.5e+09

2e+09

2.5e+09

0.5 1 1.5 2 2.5 3 3.5 4 4.5 5 5.5

RC

(nsec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(a) Task set 1

0

2e+09

4e+09

6e+09

8e+09

1e+10

1.2e+10

1.4e+10

1.6e+10

1.8e+10

2e+10

0 1 2 3 4 5 6 7 8 9 10 11

RC

(nsec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(b) Task set 2

0

2e+16

4e+16

6e+16

8e+16

1e+17

1.2e+17

0 2 4 6 8 10 12 14

RC

(nsec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(c) Task set 3

Figure 7: Task retry costs under LCM and competitor synchro-

nization methods (0.8,0.5,0.2)

0

1e+09

2e+09

3e+09

4e+09

5e+09

6e+09

7e+09

8e+09

9e+09

1e+10

0.5 1 1.5 2 2.5 3 3.5 4 4.5 5 5.5

Res-T

ime(n

sec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(a) Task set 1

0

1e+10

2e+10

3e+10

4e+10

5e+10

6e+10

0 1 2 3 4 5 6 7 8 9 10 11

Res-T

ime(n

sec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(b) Task set 2

0

5e+09

1e+10

1.5e+10

2e+10

2.5e+10

3e+10

3.5e+10

4e+10

0 2 4 6 8 10 12 14

Res-T

ime(n

sec)

Task Id

ECMRCM

LCM-EDFLCM-RMA

LF-EDF

(c) Task set 3

Figure 8: Task response times under LCM and competitor syn-

chronization methods (0.8,0.5,0.2)

446