# A Near-linear Time Constant Factor Algorithm for Unsplittable Flow Problem on Line with Bag Constraints

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1 A Near-linear Time Constant Factor Algorithm for Unsplittable Flow Problem on Line with Bag Constraints Venkatesan T. Chakaravarthy, Anamitra R. Choudhury, and Yogish Sabharwal IBM Research - India, New Delhi {vechakra, anamchou, Abstract Consider a scenario where we need to schedule a set of jobs on a system offering some resource (such as electrical power or communication bandwidth), which we shall refer to as bandwidth. Each job consists of a set (or bag) of job instances. For each job instance, the input specifies the start time, finish time, bandwidth requirement and profit. The bandwidth offered by the system varies at different points of time and is specified as part of the input. A feasible solution is to choose a subset of instances such that at any point of time, the sum of bandwidth requirements of the chosen instances does not exceed the bandwidth available at that point of time, and furthermore, at most one instance is picked from each job. The goal is to find a maximum profit feasible solution. We study this problem under a natural assumption called the no-bottleneck assumption (NBA), wherein the bandwidth requirement of any job instance is at most the minimum bandwidth available. We present a simple, near-linear time constant factor approximation algorithm for this problem, under NBA. When each job consists of only one job instance, the above problem is the same as the wellstudied unsplittable flow problem (UFP) on lines. A constant factor approximation algorithm is known for the UFP on line, under NBA. Our result leads to an alternative constant factor approximation algorithm for this problem. Though the approximation ratio achieved by our algorithm is inferior, it is much simpler, deterministic and faster in comparison to the existing algorithms. Our algorithm runs in near-linear time (O(n log 2 n)), whereas the running time of the known algorithms is a high order polynomial. The core idea behind our algorithm is a reduction from the varying bandwidth case to the easier uniform bandwidth case, using a technique that we call slicing ACM Subject Classification F.2.2 [Analysis of Algorithms] Keywords and phrases Approximation Algorithms; Scheduling; Resource Allocation Digital Object Identifier /LIPIcs.FSTTCS Introduction We consider a general resource allocation problem in which we need to schedule jobs on a system offering a certain amount of some resource (such as electrical power, processing nodes, communication bandwidth). Each job consists of a set (or bag) of job instances, out of which at most one can be chosen. Each job instance requires a particular amount of the resource for its execution. The total amount of the resource offered by the system is different at different points of time. Our goal is to choose a subset of job instances such that at any timeslot, the total amount of resource requirement does not exceed the total amount of the resource available at that timeslot. We wish to maximize the profit of the Venkatesan T. Chakaravarthy, Anamitra R. Choudhury and Yogish Sabharwal; licensed under Creative Commons License NC-ND IARCS Int l Conference on Foundations of Software Technology and Theoretical Computer Science (FSTTCS 2010). Editors: Kamal Lodaya, Meena Mahajan; pp Leibniz International Proceedings in Informatics Schloss Dagstuhl Leibniz-Zentrum für Informatik, Dagstuhl Publishing, Germany

2 182 A Near-linear Time Constant Factor Algorithm for UFP on Line with Bags chosen subset of jobs. The problem formulation is motivated by its applications in cloud computing, bandwidth allocation in networks, smart energy management and allocating processor nodes on a multi-processor system. We refer to [4, 11, 1, 12, 8] for a discussion on some of these applications. Motivated by such scheduling and bandwidth allocation scenarios, we study an abstract problem called the Varying bandwidth resource allocation problem with bag constraints (BagVBRap), introduced in [8]. We use bandwidth as a generic term to refer to the resource under contention. The BagVBRap problem generalizes several previously studied scheduling and resource allocation problems. We next define the problem formally. 1.1 BagVBRap Problem Definition The input consists of a set of jobs J. Each job J J consists of a set of job instances of which at most one can be selected for execution. A instance u of a job J is specified by an interval I u = [a, b], where a and b are the start time and the finish time of the instance u; we assume that a and b are integers. The job instance u is also associated with a bandwidth requirement ρ u and a profit p u. Let D be the maximum finish time over all instances so that the interval associated with every job instance is contained in the span [1, D]. We refer to each integer 1 t D as a timeslot. For each timeslot t, the input specifies a number B t which is the bandwidth available at timeslot t. We use the term instance as a shorthand for job instance. Let U denote the set of all n instances over all the jobs in J. For each instance u U, we view each interval I u = [a, b] as a set of timeslots in the range [a, b]. We view each job as a bag of its instances. We say that the instance u is active at a timeslot t, if t I u. For a timeslot t, let A(t) denote the set of all instances active at timeslot t. A feasible solution is a subset of instances S U such that at every timeslot t, the sum of the bandwidth requirements of the instances from S active at time t is at most B t, i.e, for every timeslot 1 t D, u S A(t) ρ u B t. We call this the bandwidth constraint. Furthermore, it is required that at most one instance is picked from each job; we call this the bag constraint; we view a job as a bag of instances and hence the terminology. The problem is to find a feasible solution S such that the sum of the profits of the jobs in S is maximized. This completes the problem description. The concept of bag constraints is quite powerful. Apart from handling the notion of release time and deadline, it can also work in a more general setting where a job can specify a set of possible time intervals where it can be scheduled. Moreover, BagVBRap allows for different instances of the same job to have different bandwidth requirements, processing times and profits. The maximum and minimum available bandwidths over all timeslots will be of use in our discussion. We denote these by B max and B min, that is B max = max t [1,D] B t and B min = min t [1,D] B t. Similarly, the maximum and minimum bandwidth requirement over all instances is also of interest. We denote these by ρ max and ρ min. That is, ρ max = max u U ρ u and ρ min = min u U ρ u. Remark: We can assume that for each instance u, the start time and end time of u are in the range [1, 2n], since this leaves the problem combinatorially unchanged. Thus, we can assume that D 2n.

3 Chakaravarthy et al Prior Work The BagVBRap problem is a generalized formulation that captures as special cases many well-studied scheduling and resource allocation problems. Here we shall describe some important special cases and then present a brief survey of some of the prior work dealing with these problems. Uniform bandwidth resource allocation problem (UBRap): This is the special case of the BagVBRap problem, where the bandwidth available is uniform across all timeslots and the bag constraints do not exist. Meaning, each job consists of only one instance and for all 1 t D, B t = B for some fixed B given as part of the input. Uniform bandwidth resource allocation problem with bag constraints (BagUBRap): This is the special case of the BagVBRap problem, where the bandwidth available is uniform across all timeslots. Varying bandwidth resource allocation problem (VBRap): This is the special case of the BagVBRap problem, where each job has only one instance. The VBRap problem is the same as the unsplittable flow problem (UFP) on line graphs, a well-studied problem. Calinescu et al. [7] presented a 3-approximation for the UBRap problem, based on LP rounding technique that they refer to as listing algorithm". Independently, Bar-Noy et al. [4] also presented a local-ratio based 3-approximation algorithm for the same problem. They also show how to handle bag constraints and derive a 5-approximation algorithm for the BagUBRap problem. A further special case of the BagUBRap problem is obtained when the bandwidth requirements and the bandwidth available are all unit. This special case has been studied under the name weighted job interval selection problem (WJISP), for which Bar-Noy et al. [5] and, independently, Berman and Dasgupta [6] presented local-ratio based 2-approximation algorithms. For the VBRap problem (i.e., the UFP problem on line) Chakrabarti et al. [9] presented an algorithm with an approximation ratio of O(log(ρ max /ρ min )). Bansal et al. [3] presented an O(log n)-approximation algorithm for the same problem. No polynomial time constant factor approximation algorithm is known for this problem. However, in a break-through result, Bansal et al. [2] obtained a quasi-ptas for UFP on line. For the BagVBRap problem, an O(log(B max /B min ))-approximation algorithm was obtained in [8]; this was achieved by extending the LP based listing algorithm of Calinescu et al. [7], Obtaining a polynomial time constant factor approximation algorithm for the VBRap problem has remained a challenging open problem. However, this has been achieved under a reasonable assumption known as the no-bottleneck assumption (NBA). No Bottleneck Assumption (NBA): We say that an input to the BagVBRap problem satisfies the no bottleneck assumption (NBA), if the maximum bandwidth requirement of every job instance is less than the minimum bandwidth available. That is, ρ max B min. Chakrabarti et al. [9] obtained the first constant factor approximation algorithm for the VBRap problem, under NBA. For the same special case, Chekuri et al. [10] improved the constant factor to (2 + ɛ). Both these algorithms are based on randomized rounding of LP solutions. 1.3 Our Result and Discussion Our main result is a constant factor approximation algorithm for the BagVBRap problem, under NBA. The running time of the algorithm is O(n log 2 n), where n is the number of job instances. The approximation ratio is 120. F S T T C S

4 184 A Near-linear Time Constant Factor Algorithm for UFP on Line with Bags An important feature of our approach is the simplicity of both the algorithm and analysis. We show how to handle the non-uniformity or the varying nature of the bandwidths available, via a reduction to the easier case of uniform bandwidths. Namely, we present a simple reduction from the BagVBRap problem with NBA to the BagUBRap problem. Given a BagVBRap input with n job instances, our algorithm produces a BagUBRap instance having at most O(n log n) job instances. We then invoke the known constant factor approximation algorithm for BagUBRap, due to Bar-Noy et al. [4], which is based on the local ratio technique, and runs in time O(n log n), where n is the number of input job instances. Thus, our algorithm for BagVBRap runs in time O(n log 2 n). Our result yields a constant factor approximation algorithm for the UFP problem on line, with NBA. The approximation ratio is 120. As mentioned earlier, Chakrabarti et al. [9] and Chekuri et al. [10] have presented constant factor approximation algorithms for this problem. The algorithm of Chekuri et al. guarantees an approximation ratio of (2 + ɛ) (for any ɛ > 0); this algorithm is based on randomized LP rounding. This algorithm offers a tradeoff between approximation ratio and running time. The best running time achievable within this framework is more than O(n 10 ). As the approximatio ratio approaches 2, the running time grows substantially. Though our algorithm is inferior in terms of approximation ratio, it is simpler, deterministic and runs in near-linear time (O(n log 2 n)). We believe that our technique of reducing varying bandwidths to uniform bandwidths may find application in other scenarios as well. Organization. The rest of the paper is devoted towards presenting our constant factor approximation algorithm for the BagVBRap problem with NBA. For a better exposition of our main ideas, in Section 3, we first consider a special case where the available bandwidths are all integral multiples of B min. For this case, we present a constant factor approximation algorithm. In Section 4, we will take care of the technicalities in dealing with the general case and derive a constant factor approximation algorithm for the general case, which runs in time O(n log 2 n). Both the algorithms use reductions to BagUBRap. 2 Preliminaries Here, we develop some notations used throughout the paper. For a set of instances X U and a timeslot t [1, D], let ρ X (t) denote the sum of bandwidth requirements of all instances in X that are active at timeslot t, i.e., ρ X (t) = u X A(t) ρ u. By a bandwidth profile P, we mean a function that specifies a bandwidth for each timeslot. If P and Q are two bandwidth profiles, we define P Q to be the profile R given by: R(t) = max{0, P (t) Q(t)}, for all timeslots t. For a bandwidth profile P and a constant c, we define c P to be the profile P given by: P (t) = c P (t), for all timeslots t. By a uniform strip of bandwidth b, we mean a profile which uniformly takes the bandwidth to be b, for all timeslots. We say that the profile P is an integral profile, if for all timeslots t, P (t) is an integral multiple of B min (we allow P (t) = 0). We say that a set of instances X fits into a profile P, if for all timeslots t, ρ X (t) P (t). The input specifies a profile P in given by: P in (t) = B t, for all timeslots t. We shall refer to P in as the input profile. Note that a feasible solution is simply a set of instances S that fits into the input profile P in and satisfies the bag constraints.

5 Chakaravarthy et al Bmin 5.Bmin 4.Bmin 3.Bmin 2.Bmin Bmin Figure 1 Illustration of slicing 3 BagVBRap with NBA: The Case of Integral Profiles In this section, for a simple exposition of our main ideas, we focus on the special case where the input bandwidth profile P in is integral. For this case, we present a constant factor approximation algorithm having running time O(n 2 log n). In the next section, we will take care of the technicalities in dealing with the non-integral input profiles and derive a constant factor approximation algorithm for the general case, which runs in time O(n log 2 n). Let K = B max /B min ; since the profile is assumed to be integral, K is an integer. We imagine the input bandwidth profile P in to be a curve giving a value for each timeslot t. We slice the area under the curve horizontally into K slices each of height B min, as illustrated in Figure 1. For 1 j K, the bandwidth profile of the jth slice is defined as follows: for 1 t D, Slice j (t) = { Bmin if B t j B min 0 otherwise Consider a feasible solution S U (which fits into the profile P in and satisfies the bag constraints). The solution S is said to be slice-respecting, if it is possible to assign the instances of S to the K slices satisfying the following two properties: (i) each instance u S is assigned to one of the K slices; (ii) for 1 j K, the subset of instances assigned to the jth slice fits into the profile Slice j. Let S opt U be the optimum solution. Our algorithm and analysis have two main components: First, we will show that the optimum solution S opt can be partitioned into 16 subsets such that each subset is a slice-respecting solution (see Section 3.1). Second, we will present an algorithm that will output a slice-respecting solution S such that the profit of S is a 5-approximation to the optimum slice-respecting solution (namely, the maximum profit solution among all slice respecting solutions). This algorithm is obtained via a reduction to the BagUBRap problem. (see Section 3.2). It follows that S is a 80-approximation to S opt. This yields the following theorem. Theorem 1. There exists a 80-approximation algorithm for the BagVBRap problem with NBA when the input bandwidth profile P in is integral. The running time of the algorithm is O(n 2 log n). 3.1 Partitioning S opt into Slice-respecting Solutions In this section, we will prove the following lemma. F S T T C S

6 186 A Near-linear Time Constant Factor Algorithm for UFP on Line with Bags Lemma 2. Any feasible solution S can be partitioned into 16 subsets such that each subset is a slice-respecting solution. The intuitive idea behind the above claim is as follows. We will delete the first slice from the bottom of the profile and obtain a new profile. We shall identify a subset of instances Y and delete them such that the remaining instances fits into the new profile. We will show that the deleted subset of instances Y can be partitioned into a collection of 16 subsets such that each subset in the collection fits into the deleted (first) slice. We shall then apply the above procedure recursively K times obtaining K such collections. A slice-respecting solution can be formed by picking one subset from each collection. This way, we can form 16 slice respecting solutions. The rest of the section is devoted to proving Lemma 2 formally. The following two lemmas are useful for this purpose. Lemma 3. Let P be any integral profile. Let X U be any subset of instances that fits into the profile P. Then, there exists a subset Y X such that: (i) for all timeslots t, ρ Y (t) min{ρ X (t), B min }; (ii) Y fits into the uniform strip of bandwidth 4 B min. Proof. For a subset T U, a job instance is said to be critical for T if removal of the job instance from T causes the total bandwidth of the remaining jobs in T to fall below B min at some timeslot. More formally, a job instance u is said to be critical for T, if ρ T \{u} (t) < B min for some t I u. We start with Y = X and then repeatedly remove jobs that are not critical for Y until no more such jobs exist. We argue that the remaining jobs in Y satisfy the required properties. The first property follows from the fact that we never remove a critical job. The second property is proved by contradiction. Suppose that for some timeslot t, ρ Y ( t) > 4 B min. Let t l = max{t t and ρ Y (t) < 2 B min } and t r = min{t t and ρ Y (t) < 2 B min }. t t First, let us suppose that both t l and t r exist. Note that by definition of t l and t r, we have that ρ Y (t) 2 B min, for all t (t l, t r ). We will now argue that there exists some job instance u Y, such that I u is contained in (t l, t r ). If this were not so, then for any timeslot t (t l, t r ), any job instance u active at timeslot t would also be active at timeslot t l or at timeslot t r. This implies that for any t (t l, t r ), ρ Y (t ) ρ Y (t l ) + ρ Y (t r ) < 4 B min. This would contradict our hypothesis that ρ Y ( t) > 4 B min. Therefore, there exists a job instance u such that I u is contained in (t l, t r ). This combined with the fact that ρ Y (t) 2 B min, for all t (t l, t r ) and the NBA implies that u is not a critical job. This contradicts our construction of Y. Now, let us assume that t l does not exist. Then ρ Y (t) 2 B min, for all t [1, t]. Let u Y be the job with the earliest finish time and let its finish time be t u. Note that ρ Y (t) 2 B min for all t [1, t u ]. This fact along with the NBA assumption implies that u is not a critical job. This contradicts our construction of Y. The case when t r does not exist is handled similar to the case of t l not existing as above. This completes the proof of the lemma. Lemma 4. Let X U be any subset of instances that fits into the uniform strip of bandwidth 4 B min. Then, X can be partitioned into a collection of 16 subsets {X 1, X 2,..., X 16 } such that each subset X i fits into the uniform strip of bandwidth B min. Proof. We say that an instance u X is large, if ρ u > B min /2; otherwise, u is said to be small. Let X l X be the set of large instances and let X s be the set of small instances. Let

7 Chakaravarthy et al. 187 us first focus on the set X l. Create 8 uniform strips of bandwidth B min, named P 1, P 2,..., P 8, called buckets. Arrange the instances in X l in a list in increasing order of their start times. Scan this list and for each instance u, add u to a bucket where it can fit without violating the bandwidth constraint. The crucial claim is that each instance u will fit into some bucket. To see this claim, consider the first instance u that does not fit into any bucket. Let the starting timeslot of u be t 0. Since we are dealing with large instances, each bucket contains exactly one large instance active at the timeslot t 0. Let Y be the set of these instances. Their combined bandwidth at t 0 is ρ Y (t 0 ) > 4 B min. This is a contradiction since X is assumed to fit into a uniform strip of bandwidth 4 B min. The argument for small tasks is similar. Consider 8 uniform strips of bandwidth B min, Q 1, Q 2,..., Q 8, which are referred to as buckets. Scan the small instances in the increasing order of their start times. For each instance u, add u to a bucket where it would fit without violating the bandwidth constraint. As before, we claim that every instance will fit into at least one bucket. To see this, suppose u be the first instance that does not fit into any bucket. Let the starting timeslot of u be t 0. For each 1 i 8, let Y i be the set of instances in Q i active at timeslot t 0. Since u is small, ρ Yi (t 0 ) > B min /2. Let Y be the union of Y 1, Y 2,..., Y 8. Their combined bandwidth at t 0 is ρ Y (t 0 ) > 4 B min. This is a contradiction since X is assumed to fit into a uniform strip of bandwidth 4 B min. The set of instances added to the 16 buckets P 1, P 2,..., P 8 and Q 1, Q 2,..., Q 8 are taken to be the required subsets X 1, X 2,..., X 16. We now prove Lemma 2. Let P 0 = P in be the input profile and let X 0 = S be the given solution. We first invoke Lemma 3 with P 0 and X 0 as inputs and obtain a set of instances Y 1. We then apply Lemma 4 to partition Y 1 into a collection of 16 subsets {Y1 1, Y1 2,..., Y1 16 }. Note that each of these subsets Y1 i fits into the profile of the first slice. We delete the first slice (bottom-most slice) from the profile P 0 and obtain a profile P 1 (formally, we set P 1 = P 0 Slice 1 ). We also delete the set of instances Y 1 from X 0 and get a new set of instances X 1 = X 0 Y 1. Observe that the set of instances X 1 fits into the profile P 1. This allows us to apply the above process starting with P 1 and X 1 as inputs. Overall, we will apply the above process iteratively K times. Formally, for j = 1 to K, we do as follows: Invoke Lemma 3 with P j 1 and X j 1 as inputs and obtain a set of instances Y j. Invoke Lemma 4 to partition Y j into a collection of 16 subsets {Yj 1, Y j 2,..., Y j 16 }. Note that each set Yj i fits into the profile of the jth slice. Define a new integral profile P j = P j 1 Slice j. Define X j = X j 1 Y j. The set of instances X j fits into the profile P j. We now form 16 solutions: for 1 i 16, let Z i = K j=1 Y i j. Each set Z i is a slicerespecting solution. This completes the proof of Lemma Approximating the Optimum Slice-respecting Solution In this section, we shall present an algorithm for finding a 5-approximation to the optimum slice-respecting solution. We achieve this goal via a reduction to the BagUBRap problem, for which Bar-Noy et al. [4] designed a local-ratio based 5-approximation algorithm. Let S be the optimum slice-respecting solution. Our goal is to find a 5-approximation to S, via a reduction to the BagUBRap problem. Let I be the input instance, which we will transform into an input instance Ĩ of BagUBRap. Let the timespan of I be [1, D] so that all the job instances finish before D. In the transformed instance Ĩ, the timespan will be [1, KD], where K is the number of slices. The bandwidth profile of Ĩ is obtained by concatenating F S T T C S

8 188 A Near-linear Time Constant Factor Algorithm for UFP on Line with Bags the bandwidth profiles of the K slices; namely, for 1 j K, the range [1 + (j 1)D, jd] corresponds to the jth slice. Formally, in the input instance Ĩ, the bandwidth available B t is declared as follows: for 1 j K and for each timeslot t [1 + (j 1)D, jd], we set B t = Slice j (t (j 1)D). Note that in Ĩ, the bandwidths available at all timeslots are either B min or 0. For each job J J of the input I, we create a corresponding job J in Ĩ. For each job instance u U in the input instance I with associated interval I u = [a, b], we create at most K copies of u as follows. For each slice 1 j K, we create a copy of u, if the instance u can be scheduled in the jth slice (meaning, for all t [a, b], Slice j (t) 0); this copy of u is declared to have the interval Iu j = [(j 1)D + a, (j 1)D + b]. This way, at most K copies are created for the instance u. Each of these copies will have the same bandwidth requirement and profit as the instance u. If J J is the job to which the instance u belongs, then all these copies of u are made instances of the corresponding job J in the transformed input Ĩ. For instance, if a job J J had d instances in I, its corresponding job in Ĩ would have at most dk instances. A minor issue in the above construction is that in Ĩ, the bandwidths available are not uniform across all the timeslots. However, these are all either B min or 0. This is easily addressed via compressing the profile by deleting all timeslots where the bandwidth availability is 0. This way, we get a BagUBRap instance. This completes the reduction. It is easy to see that a slice-respecting feasible solution S to I can be converted into a feasible solution S for Ĩ and vice versa. To see the forward direction, if u is an instance picked in the solution S and assigned to slice j, we pick the instance Iu j in S. For the reverse direction, if the instance Iu j is picked in the solution S, we include the instance I u in S and assign it to the slice j. Now invoking the 5-approximation algorithm for the BagUBRap problem [4], we get the desired 5-approximation to the optimum slice-respecting solution. The 5-approximation algorithm for the BagUBRap problem runs in time O(ñ log(ñ)), where ñ is the number of job instances in Ĩ. In our reduction, for each job instance u I, at most K instances are created in Ĩ and hence, ñ Kn. Recall that K = B max/b min. We can assume that B max nρ max. Under NBA, we also have that B min ρ max. Hence, K = O(n). Thus, our algorithm runs in time O(n 2 log n). 4 A Constant Factor Approximation for the General Case We now consider the general case involving non-integral input profiles and derive a constant factor approximation algorithm having running time O(n log 2 n). The core idea of slicing is the same as that of the earlier section. The new result is obtained by dealing with some technicalities. First, instead of slicing the input profile into slices of equal height, here we shall slice the profile into slices of geometrically increasing heights. This enables us to bring down the running time. Secondly, we shall floor" the input profile by carefully decreasing the bandwidth available at every timeslot so that it becomes geometric" and it is therefore, suitable for geometric slicing". However, the flooring" increases the approximation ratio by a constant factor. We say that a profile P is a geometric integral profile (GIP), if one of the following two conditions is true: For all timeslots t, either P (t) = 0 or P (t) = 2 i B min, for some integer i 0. There exists an integer constant c 0 such that for all timeslots t, either P (t) = 0 or P (t) = (2 i 2 c )B min, for some integer i 0. In the former case the profile is said to be a type-1 GIP and in the latter case, the profile

9 Chakaravarthy et al Bmin 2.Bmin Bmin Figure 2 Illustration of slicing said to be a type-2 GIP. The idea of having two different types of GIPs is that we shall start with a type-1 GIP and as we delete slices iteratively, we obtain type-2 GIPs. We first convert the input profile P in into a type-1 GIP by flooring the bandwidths, as follows. Define a new profile P in given by P in (t) = 2 i B min, where i = log(p in (t)/b min ). Figure 2 presents the flooring of the profile shown in Figure 1. A feasible solution S U (which fits into the profile P in ) may not fit into the profile P in. Our algorithm will actually output a solution that fits into the profile P in. In this process, we lose a constant factor in the approximation. Similar to the previous section, we now define the notion of slices and slice-respecting solutions. Let K = 1 + log(b max /B min ). We slice the profile P in horizontally into K slices in a geometric manner. The profile of the first slice Slice 1 is the uniform strip of bandwidth B min. For 2 j K, the bandwidth profile of the jth slice is given by the following profile: for 1 t D, Slice j (t) = { 2 j 2 B min if P in (t) 2 j 1 B min 0 otherwise See Figure 2 for an illustration of the slices. Consider a feasible solution S U (which fits into the profile P in and satisfies the bag constraints). The solution S is said to be slice-respecting, if it is possible to assign the instances of S to the K slices satisfying the following two properties: (i) each instance u S is assigned to one of the K slices; (ii) for 1 j K, the subset of instances assigned to the jth slice fits into the profile Slice j. Let S opt U be the optimum solution. Our algorithm and analysis have two main components: First, we will show that the solution S opt can be partitioned into 24 subsets such that each subset is a slice-respecting solution. Second, we will present an algorithm that will output a slice-respecting solution S such that the profit of S is a 5-approximation to the optimum slice-respecting solution. It follows that S is a 120-approximation to S opt. This yields the following theorem. Theorem 5. There exists a 120-approximation algorithm for the BagVBRap problem with NBA. The running time of the algorithm is O(n log 2 n). We now focus on the first component and prove a lemma similar in spirit to Lemma 2. Lemma 6. Any feasible solution S U can be partitioned into 24 subsets such that each subset is a slice-respecting solution. F S T T C S

10 190 A Near-linear Time Constant Factor Algorithm for UFP on Line with Bags We first state two variants of Lemma 3 that deal with the two types of GIPs. The proofs are similar to that of Lemma 3. Lemma 7. Let P be any type-1 GIP. Let X U be any subset of instances that fits into the profile 2 P. Then, there exists Y X such that: (i) for all timeslots t, ρ Y (t) min{ρ X (t), 2 B min }; (ii) Y fits the uniform strip of bandwidth 6 B min. Lemma 8. Let P be any type-2 GIP with parameter c 0. Let X U be any subset of instances that fits into the profile 2 P. Then, there exists Y X such that: (i) for all timeslots t, ρ Y (t) min{ρ X (t), 2 2 c B min }; (ii) Y fits into the uniform strip of bandwidth 6 2 c B min. We next state a variant of Lemma 4. The proof is similar to that of Lemma 4. Lemma 9. Let X U be any subset of instances that fits into the uniform strip of bandwidth 6 α B min, for some integer α 1. Then, X can be partitioned into 24 subsets X 1, X 2,..., X 24 such that each subset X i fits into the uniform strip of bandwidth α B min. We now prove Lemma 6. The proof is similar to that of Lemma 2 and is proved in an iterative manner. In the first iteration, we apply Lemma 7 and Lemma 9 on profile P in, which is a type-1 GIP. In the subsequent iterations, we will apply Lemma 8 and 9, as the profiles considered in these iterations are type-2 GIPs. Let P 0 = P in be the input profile, which is a type-1 GIP. Let X 0 = S be the given solution that fits into the profile P in. Observe that X 0 fits into the profile 2 P 0. We first invoke Lemma 7 with P 0 and X 0 as inputs and obtain a set of instances Y 1. We then apply Lemma 9 to partition Y 1 into 24 subsets Y1 1, Y1 2,..., Y1 24 (here α = 1). Note that each of these subsets Y1 i fits into the profile Slice 1. We delete the first slice (bottom-most slice) from the profile P 0 and obtain a profile P 1 (formally, we set P 1 = P 0 Slice 1 ). We also delete the set of instances Y 1 from X 0 and get a new set of instances X 1 = X 0 Y 1. The profile P 1 is a type-2 GIP with parameter c = 0 and X 1 fits into 2 P 1. We invoke Lemma 8 with P 1 and X 1 as inputs and obtain a set of instances Y 2. We then apply Lemma 9 to partition Y 2 into 24 subsets Y2 1, Y2 2,..., Y2 24 (here α = 2 0 = 1). Note that each of these subsets Y2 i fits into the profile Slice 2. We now delete the second slice from the profile P 1 and obtain a profile P 2 (formally, we set P 2 = P 1 Slice 2 ). We also delete the set of instances Y 2 from X 1 and get a new set of instances X 2 = X 1 Y 2. The profile P 2 is a type-2 GIP with parameter c = 1 and X 2 fits into 2 P 2. We can now iteratively apply the above process starting with P 2 and X 2 as inputs. We continue like this for K iterations. Formally, for j = 2 to K, do as follows: Invoke Lemma 8 with P j 1 and X j 1 as inputs and obtain a set of instances Y j. Invoke Lemma 9 to partition Y j into 24 subsets Yj 1, Y j 2,..., Y j 24. Each set Yj i fits into the profile Slice j. Define a profile P j given by P j = P j 1 Slice j. The profile P j is a type-2 GIP with parameter c = j 2. Define X j = X j 1 Y j. The set of instances X j fits into the profile 2 P j. We now form 24 solutions: for 1 i 24, let Z i = K j=1 Y i j. Each set Z i is a slicerespecting solution. This completes the proof of Lemma 6. We now outline the second component required to prove Theorem 5. What we need is an algorithm that approximates the optimum slice respecting solution within a factor of 5. The construction is the same as that of Section 3.2. However, we do not get a BagUBRap instance, since the bandwidths available are not uniform. Nevertheless, we can scale the

11 Chakaravarthy et al. 191 slices (and the associated job instances), so as to make all the slices of uniform bandwidth. This way we can get a BagUBRap instance. In the current scenario K = O(log(B max /B min )). Now an analysis similar to the one performed in the previous section shows that the running time of our algorithm is O(n log 2 n). References 1 S. Albers. Resource Management in Large Networks. In J. Lerner, D. Wagner, and K. Zweig, editors, Algorithmics for Large and Complex Networks: Design, Analysis, and Simulation, pages Springer Berlin/Heidelberg, N. Bansal, A. Chakrabarti, A. Epstein, and B. Schieber. A quasi-ptas for unsplittable flow on line graphs. In ACM Symposium on Theory of Computing (STOC), pages , N. Bansal, Z. Friggstad, R. Khandekar, and M. Salavatipour. A logarithmic approximation for unsplittable flow on line graphs. In ACM-SIAM Symposium on Discrete Algorithms (SODA), pages , A. Bar-Noy, R. Bar-Yehuda, A. Freund, J. Noar, and B. Schieber. A unified approach to approximating resource allocation and scheduling. Journal of the ACM, 48(5): , A. Bar-Noy, S. Guha,, J. Noar, and B. Schieber. Approximating the throughput of multiple machines in real-time scheduling. Siam Journal of Computing, 31(2): , P. Berman and B. Dasgupta. Multi-phase algorithms for throughput maximization for real-time scheduling. Journal of Combinatorial Optimization, 4: , G. Calinescu, A. Chakrabarti, H. Karloff, and Y. Rabani. Improved approximation algorithms for resource allocation. In Prooceedings of the 9th International Conference on Interger Programming and Combinatorial Optimization, V. Chakaravarthy, V. Pandit, Y. Sabharwal, and D. Seetharam. Varying bandwidth resource allocation problem with bag constraints. In 24th IEEE International Parallel and Distributed Processing Symposium (IPDPS), A. Chakrabarti, C. Chekuri, A. Gupta, and A. Kumar. Approximation algorithms for the unsplittable flow problem. Algorithmica, 47(1):53 78, C. Chekuri, M. Mydlarz, and F. Shepherd. Multicommodity demand flow in a tree and packing integer programs. ACM Transactions on Algorithms, 3(3), T. Erlebach and F. Spieksma. Interval selection: Applications, algorithms, and lower bounds. Journal of Algorithms, 46(1):27 53, M. Flammini, G. Monaco, G. Moscardelli, H. Shachnai, M. Shalom, T. Tamir, and S. Zaks. Minimizing total busy time in parallel scheduling with application to optical networks. In 23rd IEEE International Parallel and Distributed Processing Symposium (IPDPS), F S T T C S

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