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Online Learning in Online Auctions?
Avrim Blum
Department of Computer Science, Carnegie Mellon University, Pittsburgh, PA1
Vijay Kumar
Strategic Planning and Optimization Team, Amazon.com, Seattle, WA
Atri Rudra
Department of Computer Science, University of Texas at Austin, Austin, TX2
Felix Wu
Computer Science Division, University of California at Berkeley, Berkeley, CA3
Abstract
We consider the problem of revenue maximization in online auctions, that is, auc-
tions in which bids are received and dealt with one-by-one. In this paper, we demon-
strate that results from online learning can be usefully applied in this context, and
we derive a new auction for digital goods that achieves a constant competitive
ratio with respect to the optimal (offline) fixed price revenue. This substantially
improves upon the best previously known competitive ratio for this problem of
O(exp(√loglogh)) [4]. We also apply our techniques to the related problem of de-
signing online posted price mechanisms, in which the seller declares a price for each
of a series of buyers, and each buyer either accepts or rejects the good at that price.
Despite the relative lack of information in this setting, we show that online learning
techniques can be used to obtain results for online posted price mechanisms which
are similar to those obtained for online auctions.
?Portions of this work appeared as an extended abstract in Proceedings of
SODA’03 [5].
Email addresses: avrim@cs.cmu.edu (Avrim Blum), vijayk@amazon.com
(Vijay Kumar), atri@cs.utexas.edu (Atri Rudra), felix@cs.berkeley.edu
(Felix Wu).
1Supported in part by National Science Foundation grants CCR-0105488 and IIS-
0121678.
2Work done while the author was at IBM India Research Lab, New Delhi, India.
3Supported in part by National Science Foundation ITR grant CCR-0121555.
Preprint submitted to Elsevier Science16 September 2003
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1 Introduction
Auctions are traditional and well-studied economic mechanisms, and economists
have long studied the design of auctions intended to satisfy various goals, in-
cluding that of maximizing the total revenue obtained by the auctioneer from
the auction. Traditionally, however, economists have analyzed auctions under
the assumption that statistical information about the participating bidders is
available. Recent work in computer science has been directed toward designing
auctions in the absence of such statistical assumptions, using instead a form
of worst-case competitive analysis [3,4,7,10,12].
The proliferation of Internet auctions and the increasing availability of media
on the Internet has prompted particular attention to the design of auctions for
digital goods, that is, goods available in unlimited supply [7,10]. In this paper,
we focus on such goods, though our techniques may also be useful in the case
of limited supply goods. A key property of digital goods is that it will often
be useful to conduct auctions of such goods over time, with bidders arriving
one-by-one, rather than as a group. Hence, we are interested here in designing
online auctions for digital goods, a problem first described by Bar-Yossef et
al. [4].
In the model of Bar-Yossef et al. [4], n bidders arrive in a sequence. Each
bidder i is interested in one copy of the good, and values this copy at vi. The
valuations are normalized to some range [1,h], so that h is the ratio between
the highest and lowest possible valuations. Bidder i places bid bi, and the
auction must then determine whether to sell the good to bidder i, and if so,
at what price si≤ bi. This is equivalent to determining a sales price si, such
that if si≤ bi, bidder i wins the good and pays si; otherwise, bidder i does
not win the good and pays nothing.
The utility of a bidder is then given by vi− siif bidder i wins; 0 if bidder i
does not win. As in Bar-Yossef et al. [4], we are interested in auctions which
are incentive-compatible, that is, auctions in which each bidder’s utility is
maximized by bidding truthfully and setting bi= vi. As shown in that paper,
this condition is equivalent to the condition that each si depends only on
the first i − 1 bids, and not on the ith bid. Hence, the auction mechanism is
essentially trying to guess the ith valuation, based on the first i−1 valuations.
Note that in an online auction, the sales prices siare not actually revealed to
the bidders, since we need the bidders to declare their valuations, so that the
auction can use this information in dealing with future bidders. In auctions
conducted remotely over networks, however, the bidders may not trust the
auctioneer to set sales prices before seeing the next bid. Buyers would clearly
prefer to receive these sales prices directly and then to make a decision ac-
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cordingly whether or not to purchase the good. (Buyers purchase if and only if
si≤ vi.) We call such a mechanism a posted price mechanism [11]. The trade-
off in using such a mechanism is that in exchange for the greater trust of the
buyers, the seller loses the complete information about the buyers’ valuations.
As in previous papers [4,10,12], we will use competitive analysis to analyze the
performance of any given auction or mechanism. That is, we are interested in
the worst-case ratio (over all sequences of valuations) between the revenue of
the “optimal offline” auction and the revenue of the online auction. Following
previous papers [4,10], we take the optimal offline auction to be the one which
optimally sets a single fixed price for all bidders. Thus, our goal is what is
sometimes called “static optimality.” The revenue of the optimal fixed price
auction is given by F(v) = maxi{vini}, where ni = |{j | vj≥ vi}| is the
number of bidders with valuation at least vi. An online auction A with revenue
RA(v) is said to be c-competitive if for any sequence v, RA(v) ≥ F(v)/c. We
take RAto be the expected revenue if A is randomized.
In Section 2, we present an asymptotically constant-competitive online auction
for digital goods. By asymptotically, we mean that our auction achieves a
revenue which is a constant fraction of F, but minus an additive term. (In our
case, this term is O(hlnlnh).) Hence, as F becomes large, this additive term
becomes negligible. Nevertheless, it is important to minimize this term, since it
roughly corresponds to the size of the smallest auctions for which we can give
good revenue bounds. Theorem 4 gives a general lower bound showing that
our additive constant is nearly optimal: in particular, any constant-competitive
algorithm must have an additive constant Ω(h).
In Section 3, we derive a similar result for the problem of designing online
posted price mechanisms. (Offline posted price mechanisms have been previ-
ously studied by Hartline [11].) Such mechanisms provide much less informa-
tion to the auctioneer about the bidders’ valuations, but surprisingly, we are
still able to obtain results very similar to those obtained in the online auction
setting.
Our results are based on application of machine learning techniques to the
online auction problem. Setting a single fixed price for the auction can be
thought of as following the advice of a single “expert” who predicts that fixed
price for every bidder. Performing well relative to the optimal fixed price is
then equivalent to performing well relative to the best of these experts, a prob-
lem well-studied in learning theory [2,6,8,9,13]. The posted price setting then
corresponds to a version of the “bandit” problem [2], in which the informa-
tion received depends on the expert chosen at each step. Our algorithms are
derived by adapting these techniques to the online auction setting.
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2 Online auctions: the full information game
We use a variant of Littlestone and Warmuth’s weighted majority (WM)
algorithm [13] given in Auer et al. [1,2]. In our context, let X = {x1,...,x?}
be a set of candidate fixed prices, corresponding to a set of experts. Let rk(v)
be the revenue obtained by setting the fixed price xkfor the valuation sequence
v, and let FX(v) = maxkrk(v) be the optimal fixed price revenue on sequence
v, when restricted to fixed prices in X.
Given a parameter α ∈ (0,1], define weights wk(i) = (1+α)rk(v1,...,vi)/h. Clearly,
the weights can be easily maintained using a multiplicative update. Then, for
bidder i, the auction chooses si∈ X with probability
wk(i − 1)
??
pk(i) = Pr[si= xk] =
j=1wj(i − 1).
This algorithm is shown in Figure 1.
Algorithm WM
Parameters: Reals α ∈ (0,1] and X ∈ [1,h]?.
Initialization: For each expert k, initialize rk() = 0,wk(0) = 1.
For each bidder i = 1,...,n:
Set the sales price sito be xkwith probability pk(i) =
wk(i−1)
??
j=1wj(i−1).
Observe bi= vi.
For each expert k, update rk(v1,...,vi) and wk(i) = (1 + α)rk(v1,...,vi)/h.
Fig. 1. WM in our setting
The following theorem appears in Auer et al., with the proof adapted from
proofs appearing in Freund and Schapire [9] and Littlestone and Warmuth [13].
Theorem 1 [1, Theorem 3.2] For any sequence of valuations v, the revenue
of auction WM is at least:
RWM(v) ≥ (1 −α
2)FX(v) −hln?
α
.
For completeness, we provide a proof here.
Proof. Let gk(i) denote the revenue gained by the kth expert from bidder i,
that is, gk(i) = xk, if vi≥ xk, and gk(i) = 0 otherwise. Then, rk(v1,...,vi) =
gk(i)+rk(v1,...,vi−1). Let W(i) =??
bidder i.
k=1wk(i) be the sum of the weights after
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The expected revenue of the auction from bidder i + 1 is given by
gWM(i + 1) =
??
k=1wk(i)gk(i + 1)
W(i)
.
We can then relate the change in W(i) to the expected revenue of the auction
as follows:
W(i + 1)=
?
?
?
?
k=1
wk(i)(1 + α)gk(i+1)/h
≤
k=1
wk(i)(1 + α(gk(i + 1)/h))
=W(i) + α
?
?
k=1
wk(i)(gk(i + 1)/h)
=W(i)(1 + α(gWM(i + 1)/h)),
where for the inequality, we used the fact that for x ∈ [0,1], (1+α)x≤ 1+αx.
Since W(0) = ?, we have
W(n) ≤ ? ·
n
?
i=1
(1 + α(gWM(i)/h)).
On the other hand, the sum of the final weights is at least the value of the
maximum final weight. Hence, W(n) ≥ (1 + α)FX/h.
Taking logs, we have
FX
h
ln(1 + α) ≤ ln? +
n
?
i=1
ln(1 + α(gWM(i)/h)).
For x ∈ [0,1], x −x2
FX
h
2≤ ln(1 + x) ≤ x; hence,
≤ ln? +α
?
α −α2
2
?
hRWM.
Rearranging this inequality yields the theorem.
2
Now let X consist of all powers of (1 + β) between 1 and h. If we take α =
β =?
3, we get the following theorem.
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Theorem 2 For any ? ∈ (0,1], restricting to valuation sequences with F(v) ≥
24h
?2(lnlnh + ln(4
of (1 +?
?)), auction WM with α =?
3) is (1 + ?)-competitive relative to the optimal fixed price revenue.
3and X consisting of all powers
Proof. First note that F(v) ≤ (1+β)FX(v), since rounding down to a power
of (1 + β) loses at most a factor of (1 + β) in the revenue. From Theorem 1,
we have
RWM(v) ≥ (1 −α
2)FX(v) −hln?
α
.
Note that ln(1 + β) ≥ β − β2/2 = ?/3 − ?2/18 ≥ ?/4. Hence, by construction,
hln(
≤3h
hln?
α
=
lnh
ln(1+β))
?
3
?(lnlnh + ln(4
?)) ≤?
8F(v).
Thus,
RWM(v)≥(1 −?
6)F(v)
(1 +?
3)−?
8F(v)
≥(1 −?
6−?
8(1 +?
3))F(v)
(1 +?
3)
≥(1 −?
3)F(v)
(1 +?
3)≥F(v)
1 + ?.
2
For any moderately large auction, the performance guarantee of the weighted
majority auction mechanism is dramatically better than that of previous auc-
tion mechanisms. As a comparison, Bar-Yossef et al. show that their weighted
buckets auction is O(exp(√loglogh))-competitive [4]. However, in that case,
the competitive ratio is achieved for valuation sequences with F(v) ≥ 4h. The
following theorem (Theorem 3) shows that WM fails on such small valuation
sequences, and indeed, the theorem provides a fairly tight lower bound on the
sequences for which WM succeeds in achieving a constant competitive ratio.
In Theorem 4, we then prove that any algorithm achieving a constant compet-
itive ratio must lose an additive term Ω(h) in the revenue. (Equivalently, it is
not possible to achieve a constant competitive ratio when F(v) = o(h).) Thus
there is an O(loglogh) gap in the additive term between the performance of
WM (Theorem 2 above) and our general lower bound.
Theorem 3 For any function f(h) = o(hloglogh), even when restricted
to valuation sequences with F(v) ≥ f(h), WM with any constant α is not
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constant-competitive. Furthermore, this holds even if WM is allowed to begin
with unequal initial weights.
Proof.
distinct and the initial weights are all equal (as in the algorithm described
in Figure 1). In this case, note that if the competitive ratio is at most some
constant c, then for every value x ∈ [1,h], there must be some xk∈ X such
that xk≤ x ≤ cxk. Otherwise, a sequence of bids of value x would lead to a
competitive ratio more than c. Hence, ? ≥ logch = Ω(logh).
We first prove the claim under the assumption that the xk are all
Now consider a bid sequence consisting entirely of bids of value x1 = 1. If
there are n bids, clearly F = n. For k ?= 1, for all i, wk(i) = 1, while w1(i) =
(1 + α)i/h. Hence, the expected revenue from the ith bidder is no more than
1
?(1 + α)i/h. Summing over the n bidders, we get a total revenue of at most
n
?(1+α)n/h. If the competitive ratio is at most c, then we need (1+α)n/h≥?
which implies n = Ω(hlog?) = Ω(hloglogh), from which the result follows.
c,
The above argument implicitly assumes all xi are distinct (or, equivalently,
that WM begins with all experts having the same weight). We can generalize
the lower bound to hold even when experts begin with different weights as
follows. As before, suppose the competitive ratio is at most c. Then, for any
value x ∈ [1,h], let qxbe the fraction of initial weight on experts xi∈ [x
Consider a sequence of n bids at the value x for which qxis smallest. In this
case, F = nx. The online algorithm makes at mostnx
window, and at most nxqx(1 + α)nx/hfrom experts inside this window. Since
qx≤ 1/log2ch and since c-competitiveness implies an online revenue of at least
nx
c, it must be that (1+α)nx/h≥ (log2ch)/2c and therefore nx = Ω(hloglogh).
Thus, the result again follows.
2c,x].
2cfrom experts below this
2
A bid sequence consisting entirely of bids of one value may seem somewhat
anomalous; in particular, h does not represent the true ratio between the
highest and lowest valuations, and most of the weights remain at their initial
value. However, the example does not depend on these properties. To see this,
one can prepend to the sequence above a set of bids, including a bid at h, such
that the revenue obtained from the prefix by using any fixed price xi∈ X falls
in the range [h,2h]. Since in the prefix F = O(h), for any auction, the bids in
the prefix can be ordered in such a way that the auction achieves revenue at
most O(h) from these bids.
It is not possible to do much better using some other algorithm. We show here
that any constant-competitive algorithm must lose an additive term Ω(h),
using analysis similar to that used for one-way trading.
Theorem 4 There is no constant-competitive algorithm for all valuation se-
quences with F(v) ≥ f(h) when f(h) = o(h). Equivalently, suppose A is an on-
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line algorithm such that for all valuation sequences v, RA(v) ≥ F(v)/c−f(h),
where c is a constant. Then f(h) = Ω(h).
Proof. First note that the two statements of the theorem are equivalent. In
one direction, if we have an algorithm with competitive ratio c and additive
term −f(h), then for F(v) ≥ 2cf(h), the algorithm will be 2c-competitive.
In the other direction, if we have an algorithm with competitive ratio c for
F(v) ≥ f(h), then it is (trivially) c-competitive with an additive term −f(h)
on the smaller sequences. We prove the second statement below.
Let A be an online algorithm with constant competitive ratio c and additive
term −f(h). Let k = 2c and m = 2kk−1. We will show that f(h) ≥ h/(km).
Consider the very first bid, and let Pr[a,b] denote the probability that A’s
sales price is in the range [a,b]. Suppose it is the case that Pr[1,h/m] ≤ 1/k.
Then, if the bid comes in at h/m, the online algorithm’s expected gain is at
most h/(km) but F(v) = h/m. Thus, f(h) ≥ F(v)/c − RA(v) ≥ h/(km). So,
we can assume that Pr[1,h/m] > 1/k.
In general, define the series Lt as follows: L0 = 0 and Lt+1 = h/m + kLt.
So, Lt+1 = h/m + hk/m + ... + hkt/m. By definition of m, Lk ≤ h. So,
there must be some interval (Lt,Lt+1] ⊆ [1,h] such that Pr(Lt,Lt+1] ≤ 1/k.
As above, suppose the first bid comes in at Lt+1. In this case, the online
algorithm’s expected gain is at most Lt+Lt+1/k, but F(v) = Lt+1. So, cf(h) ≥
F(v) − cRA(v) ≥ Lt+1− c(Lt+ Lt+1/k) = Lt+1/2 − cLt. Plugging in the
definition of Lt+1, this is at least h/(2m), and thus f(h) ≥ h/(km).
2
3 Posted price mechanisms: the partial information game
As noted in Section 1, the seller using an online posted price mechanism is
at a considerable disadvantage compared to a seller using an online auction,
since with a posted price mechanism, the seller receives much less information
about the buyers’ valuations. Nevertheless, as described below, it is still pos-
sible to design an online algorithm which achieves (asymptotically) a constant
competitive ratio with respect to the optimal fixed price revenue.
To do this, we use a version of the algorithm Exp3 of Auer et al. [1]. As with
an online auction, the choice of a sales price corresponds to the choice of an
expert. However, in an online auction, the subsequent bid reveals exactly how
well each expert would have done. In a posted price mechanism, at each step,
we will know what would have happened with some, but not all, of the possible
sales prices. The only sales price whose performance we are guaranteed to know
is the one chosen: this corresponds to an online learning algorithm which uses
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only information about the gain of the chosen expert at each step.
The algorithm Exp3 essentially contains algorithm WM, described in Section
2, as a subroutine. At each step, we take the probability distribution p used by
WM and mix it with the uniform distribution to obtain a modified probability
distribution p, which is then used to select an expert. Following each buyer’s
accept/reject decision, we use the information obtained about the gain of the
chosen expert to formulate a simulated gain vector, which is then used to
update the weights maintained by WM.
Figure 2 describes the algorithm Exp3 in our setting.
Algorithm Exp3
Parameters: Reals α ∈ (0,1], γ ∈ (0,1], and X ∈ [1,h]?.
Initialization: For each expert k, initialize rk(0) = 0, wk(0) = 1.
For each buyer i = 1,...,n:
Set the posted price sito be xkwith probability
pk(i) = (1 − γ)pk(i) +γ
where
pk(i) =
??
For the chosen price si= xk∗,
if buyer i accepts, set gk∗(i) = si, else set gk∗(i) = 0;
set gk∗(i) =γ
?
For all other experts k, set gk(i) = 0.
For all experts k,
update rk(i) = rk(i − 1) + gk(i) and wk(i) = (1 + α)rk(i)/h.
?,
wk(i−1)
j=1wj(i−1).
gk∗(i)
pk∗(i).
Fig. 2. Exp3 in our setting
Theorem 4.1 in Auer et al. then becomes the following:
Theorem 5 [1, Theorem 4.1] For any sequence of valuations v, the revenue
of auction Exp3 is at least:
RExp3(v) ≥ (1 − γ −α
2)FX(v) −h?ln?
αγ
.
As above, let X consist of all powers of (1+β) between 1 and h. For appropriate
choices of α, β, and γ, we get the following theorem.
Theorem 6 For any ? ∈ (0,1], restricting to valuation sequences with F(v) ≥
2304h
?4
lnh(lnlnh + ln(4
sisting of all powers of (1 +?
fixed price revenue.
?)), mechanism Exp3 with α =
3) is (1 + ?)-competitive relative to the optimal
?
6, γ =
?
12, and X con-
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Proof. As in Theorem 2, F(v) ≤ (1 + β)FX(v), and again, ln(1 + β) ≥ ?/4.
In this case, we have
h?ln?
αγ
=
h(
lnh
ln(1+β))ln(
(?
lnh
ln(1+β))
12)
6)(?
≤288h
?3 lnh(lnlnh + ln(4
?)) ≤?
8F(v).
Thus, by Theorem 5,
RExp3(v) ≥ (1 −
?
12−
?
12)F(v)
(1 +?
3)−?
8F(v) ≥
F(v)
(1 + ?),
using the same calculations as in Theorem 2.
2
Again, we can show that this mechanism is not constant-competitive on val-
uation sequences with small fixed price revenue.
Theorem 7 For any function f(h) = o(hloghloglogh), even when restricted
to valuation sequences with F(v) ≥ f(h), Exp3 with any constant α is not
constant-competitive.
Proof. Suppose the competitive ratio is at most some constant c. As before,
we must have ? = Ω(logh). Again consider a valuation sequence consisting
entirely of valuations at x1= 1, and let n denote the number of buyers, so
that F = n.
For k ?= 1, wk(i) = 1 for all i. Hence, because r1(i) is nondecreasing, w1(i),
p1(i), and p1(i) are all nondecreasing in i. Furthermore, the expected revenue
from buyer i is given by p1(i). Therefore, in order for the competitive ratio to
be c, we must have p1(n) ≥ 1/c.
From the definition of p, this implies that p1(n) ≥ 1/c. But, p1(n) is at most
1
?(1 + α)r1(n)/h, so we must have r1(n) ≥ hlog?
c.
Recall that r1(n) =?n
g1(i) is given by p1(i)[(γ/?)(1/p1(i))] = γ/?. Hence, we need n ≥ (?/γ)hlog?
Ω(h?log?) = Ω(hloghloglogh), and the theorem follows.
i=1g1(i). Furthermore, note that the expected value of
c=
2
The case of unequal initial weights can be handled analogously as in Theorem 3
above.
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4Extensions and Conclusions
Note that given any two auction mechanisms, we can achieve performance
which is within a factor of two of the best of the two auctions by simply
assigning probability 1/2 to each. By combining the weighted majority and
weighted buckets auctions of [4], we can achieve a constant competitive ratio
for valuation sequences with large F, while maintaining the O(exp(√loglogh))
competitive ratio for sequences with smaller F.
Also note that our techniques can be applied to the limited supply case, so
long as the sequence of bids can be truncated as soon as we run out of items to
sell. While this is not a standard notion in competitive analysis, it does suggest
that the weighted majority auction could perform well when the supply is not
too small and the bids are generated in some unknown, but non-adversarial,
manner. Using the standard notion of competitive ratio, Lavi and Nisan give
a lower bound of Ω(logh) for the limited supply case [12].
In this paper, we have demonstrated the power of online learning techniques
in the context of online auction problems by giving a (1 + ?)-competitive
online auction for digital goods. This auction requires valuation sequences
with slightly larger, but still quite reasonable, optimal fixed price revenues.
We have demonstrated that such a condition is necessary for our weighted
majority-based auction. We have also devised a (1 + ?)-competitive online
posted price mechanism under a similar assumption. This result is somewhat
surprising since the amount of information available to the algorithm is much
smaller in a posted-price scenario than in the standard online auction setting.
In both cases, the simplicity of the underlying algorithms suggests that these
mechanisms would be practical in a wide variety of settings.
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