1. Online Mechanism Design (Randomized Rounding on the Fly)
Piotr Krysta, University of Liverpool, UK
Berthold Vöcking, RWTH Aachen University, Germany

2. Combinatorial Auctions
m indivisible items (goods) given for sale
n potential buyers (bidders), each with a valuation function v(.) for subsets (bundles) of goods
v(.) may express complex preferences, e.g.:
complements: v(camera + battery) > v(camera) + v(battery)
substitutes: v(Apple iPhone + Samsung Galaxy) <
v(Apple iPhone) + v(Samsung Galaxy)
Goal: Partition m goods among n bidders to maximize the social welfare (SW)
Example:
m=2 bidders {A,B}
n=2 goods {x,y}
v( {x} )
v( {y} )
v( {x,y} ) A 1 5 7 B 3 6
Opt SW = 8

3. Combinatorial Auctions: Applications
Combinatorial auctions have many important applications: * Government Spectrum Auctions (UK, Germany, Sweden, USA, …) * Allocation of Airspace System Resources * Auctions for Truckload Transportation * Auctioning Bus Routes (London)

4. Combinatorial Auctions: Problem definition
|U|= m = 8
Combinatorial Auction (CA): n bidders
U = set of m items (goods)
Each e ε U available in b ≥ 1 copies (supply)
Bidder i has valuation f-n:
Meaning:
money i is willing to pay for S
Allocations:
Problem: compute allocation
maximizing social welfare:
b=2 1 2 3 4 vv 5 vv 6 vv 7 vv 8 vv

5. How are bidders represented ? (Demand oracles)
Problem:
The length of bidder’s valuation v(.) is exponential in m.
v(.) given by demand oracles Di(Ui, p):
Given item prices what is utility maximizing
subset Si Ui and its valuation v(Si) ?
Utility of bidder i for set Si:
Demand oracle is:
restricted if
unrestricted if 1 2 3 4 vv 5 vv 6 vv 7 vv 8 vv

6. Truthful mechanisms
-- deterministic mechanism for CA: A mechanism is truthful if for each bidder i, all vi , vi’ and all declarations v-I of the other bidders except bidder i: Randomized mechanism = prob. distribution over deterministic mechanisms. It is universally truthful if each of these mechanisms is truthful.

7. Truthfulness via direct characterization & on-line algs
Achieve truthfulness by serving bidders one by one in a given order, say i=1,2,…,n, and offering items at fixed (posted) prices: If  set of items offered to bidder i, define prices (indep. from i) and compute: * bundle Si := Di(Ui , pi) * payment (without knowing the valuations of bidders i+1,…,n) Arrival models: * random order of arrivals (secretary model) * arbitrary (adversarial) order of arrivals. Goal: find alloc. S in A maximizing the social welfare. We use standard on-line competitive analysis (CR = competitive ratio)

8. On-line models: standard definition & some aspects
Competitive ratio CR (of a randomized online algo.): Σ = set of all arrival sequences of n bidders with valuations for m items For σ ε Σ: S(σ) = alloc. computed by algo., opt(σ) = opt offline alloc. OBSERVE: Adversarial arrival model: If valuations v() are unbounded, then R cannot be bounded. REASON: The b bidders arriving last might have huge v()’s, s.t. copies cannot be given to any bidders that arrive before them. Thus: assume 1 ≤ vi(S) ≤ μ for every bidder i, S subset U. Random arrival model: We assume unbounded valuations. NOTE: Random arrivals used only to extract estimate of the bids’ range.

9. Our contributions: CAs + Random arrivals model
1. General v(): for any b ≥ 1 we obtain CR
(the first online result with )
2. General v(), bundles size ≤ d: for any b ≥ 1 we obtain CR
(previous O(d2) only for b=1)
3. XOS v() and b = 1: we obtain a CR
(the first online result for submodular/XOS v(.))
Previous results: comp. l.b best known u.b.
XOS v(): O(log (m)  log log (m))-apx univ. truthful offline mech. [Dobzinski ‘09]
General v(): apx truthful in exp. offline mech [Lavi, Swamy ’05]
-apx (b=1) univ. truthful [Dobzinski, Nisan, Schapira ‘05]
-apx (deterministically) truthful [Bartal, Gonen, Nisan ‘05]

10. Our contributions: CAs + Adversarial arrivals model
1. General v(): for any b ≥ 1 we obtain CR
(the first online result with )
2. General v(), bundles size ≤ d: for any b ≥ 1 we obtain CR
(previous O(d2) only for b=1)
3. XOS v() and b = 1: we obtain a CR
(the first online result for submodular/XOS v(.))
Previous results: comp. l.b best known u.b.
XOS v(): O(log (m)  log log (m))-apx univ. truthful offline mech. [Dobzinski ‘09]
General v(): apx truthful in exp. offline mech [Lavi, Swamy ’05]
-apx (b=1) univ. truthful [Dobzinski, Nisan, Schapira ‘05]
-apx (deterministically) truthful [Bartal, Gonen, Nisan ‘05]

11. Your most profitable bundle ?
Overselling Multiplicative Price Update (MPU) Algorithm [inspired by BGN 05, …]
b=2
Bidders:
Bidder 1 vv
Your most profitable bundle ? vv vv vv vv

12. Your most profitable bundle ?
Overselling Multiplicative Price Update (MPU) Algorithm [inspired by BGN 05, …]
b=2
Bidders:
Bidder 2 vv
Your most profitable bundle ? vv vv vv vv

13. Your most profitable bundle ?
Overselling Multiplicative Price Update (MPU) Algorithm [inspired by BGN 05, …]
b=2
Bidders:
Bidder 3 vv
Your most profitable bundle ? vv vv vv vv

14. Warm-up: Overselling Multiplicative Price Update (MPU) Algorithm [inspired by BGN 05, …]
Order of bidders 1,2,…,n is arbitrary (adversarial).
1. For each good
2. For each bidder
Set
Update for each good
NOTE: Bidder i gets set

15. Overselling MPU Algorithm: Analysis
LEMMA 1. For any : * S assigns ≤ sb copies of each item, * LEMMA 2. For : THEOREM 1. S  infeasible alloc. if :
overselling
factor
infeasible

16. Overselling MPU Algorithm: Analysis
LEMMA 1 (Part II). For any ,
where
LEMMA 2. For ,
THEOREM 1. The algorithm with outputs an infeasible alloc. S:
where copies of each item is assigned;
if , then
PROOF: (1) is by LEMMA 1 (P. I). By L. 1 (P. II):
which with LEMMA 2 gives:
By v(opt) ≥ L, we have the following and this implies claim (2): ☐

17. Overselling MPU Algorithm: Analysis
LEMMA 1 (Part I). For any , alloc. S assigns ≤ sb copies of each item, where PROOF: Consider e ε U. Suppose, after some step, copies of e assigned to bidders. price of e ≥ After this step, the algorithm might give further copies of e to bidders whose maximum valuation exceeds μL. By definition of μ, L there is ≤ 1 such bidder that receives ≤ 1 copy of e. Hence, at most copies of e assigned. ☐

18. Overselling MPU Algorithm: Analysis
LEMMA 2. For , PROOF: = feasible allocation (allocates ≤ b of each item) Algo. uses demand oracle: , so By using and summing (*) for all bidders we obtain (last “≥” follows because T allocates ≤ b copies of each item) Taking T = opt implies the claim. ☐

19. Overselling MPU Algorithm with Oblivious RR
Larger price update factor
 “more feasible” solution + worse approximation
Smaller
 helps “learn” correct prices, but, produces in-feasible solution.
Idea: Achieve feasibility and good approximation by defining appropriate sets Ui for demand oracles, and using RR.
Idea: Provisionally assign bundles Si of virtual copies to bidders following MPU algorithm
 learn correct prices
Number of virtual copies ≤ b  log(μbm) (LEMMA 1)
Oblivious randomized rounding (RR)  used to decide (with small Pr = q) which provisional bundles Si become final bundles.

20. Overselling MPU Algo. with Oblivious Randomized Rounding
Bidders:
Bidder 1 vv
Your most profitable bundle ? vv vv
 YES! (Pr=q) vv
b=2 vv

21. Overselling MPU Algo. with Oblivious Randomized Rounding
Bidders:
Bidder 2 vv
Your most profitable bundle ? vv vv
 YES! (Pr=q) vv
b=2 vv

22. Overselling MPU Algo. with Oblivious Randomized Rounding
Bidders:
Bidder 3
Your most profitable bundle ? vv vv
 NO! (Pr=1-q) vv
b=2 vv

23. MPU Algorithm with Oblivious RR
Order of bidders 1,2,…,n is arbitrary (adversarial).
1. For each good
2. For each bidder
Set
Update for each good
With prob.
Update for each good

24. MPU Algorithm with Oblivious RR: remarks
1. For each good
2. For each bidder
Set
Update for each good
With prob.
Update for each good
The algorithm outputs allocation R; payment for Ri is
Def. of Ui in line 3. ensures that R is feasible!
If q=0, then the provisional alloc. S is same as MPU algo. with Ui=U.
If q=0, then the output alloc. R is empty.
With prob. 1-q the algo. increases prices of e in Si but does not sell Si (and thus “learns” the correct prices).
If q>0, then LEMMA 1 holds, but LEMMA 2 doesn’t!
We will show a stochastic version of LEMMA 2 to imply O(1/q)-apx.

25. Overselling MPU Algo. with RR: Analysis
Recall the previous analysis: LEMMA 1. For any : * S assigns ≤ sb copies of each item, * LEMMA 2. For :
Always holds
Not always holds!!!
opt bundle for i

26. A stochastic LEMMA 2’ for CAs with d-bundles
LEMMA 2’. Consider CA with |bundles| ≤ d, and let
Then for any and any bundle :
and
THEOREM 2. The MPU algorithm with oblivious RR and q as above is
for CA with |bundles| ≤ d and multiplicity b.

27. Summary and further questions?
We design the first online (universally truthful) mechanisms achieving competitiveness for any supply b ≥ 1. New technique: we combine the online allocation of bidders with the concept of oblivious randomized rounding. Our mechanisms are simple and intuitive: each bidder’s demand oracle is queried only once, … We achieve competitive ratios close to or even beating the best known approx. factors for the corresponding offline setting. Question: The main open problem is to design similar deterministic mechanisms.

28. Thanks! Questions ?

29. Problem definition: Submodular and XOS valuations
We also consider special valuations  Submodular (decreasing marginal utilities):  XOS (fractionally subadditive): FACT: If v() if submodular then it is XOS.

30. PROOFS

31. Overselling MPU Algorithm: Analysis
LEMMA 1 (Part II). For any , where and PROOF: Let and As bidders are individually rational: , hence: Now and imply the claim.☐

32. MPU Algo. with RR: stochastic LEMMA 2’
LEMMA 2’. If for any and any then PROOF: Fix bidder i and  feasible opt alloc. By for any coin flips of the algorithm
by (**)

33. MPU Algo. with RR: stochastic LEMMA 2’
LEMMA 2’. If for any and any
then
PROOF: Sum (***) for all bidders:

34. MPU Algo. with RR: stochastic LEMMA 2’
LEMMA 2’. If for any and any
then
PROOF:
By LEMMA 1:
Now: E[v(Ri)]=qE[v(Si)] as Pr[Ri=Si]=q, so E[v(R)]=qE[v(S)], and finally E[v(R)] ≥ q  v(opt)/8. ☐

35. Proving (**) for d-bundles
LEMMA. Consider CA with |bundles| ≤ d ≥ 1, and let
Then for any and any bundle of at most d items:
PROOF: Fix bidder i. By LEMMA 1, is in
of the provisional bundles , and each of them becomes final with prob
Consider and note that if was sold times, i.e., at most b-1 of its provisional bundles became final.
Thus the prob. that is:
By and union bound ☐

36. END