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Explicit Aggregation

To handle large numbers of names which appear globally in name-based routing tables, NBRP supports combining collections of name suffixes that map to the same routing information into routing aggregates. For instance, we expect an ISP content router to group all of the names from its customers into a small number of aggregates. Routing updates then consist of a small number of aggregates rather than the large number of individual name entries contained in each aggregate. Load on an entire data center or network may be advertised as load on the aggregates advertised by that data center or network.

Routing aggregate advertisements contain a version number, so that a content router can detect a change in the contents of an aggregate. Aggregate contents are discovered by sending an INRP request back to the router advertising the aggregate; this request is for a ``diff'' between the last known version and the advertised version, so that large aggregates do not have to be resent in their entirety. Aggregate membership is relatively long-lived, compared to dynamic routing state, so that content routers can amortize the cost of learning the names in an aggregate over many routing updates.

All names in a routing aggregate are treated identically in routing calculation, thus reducing load at content routers. This is accomplished in our implementation by mirroring the indirection provided by aggregates in the routing table, as shown in figure 5. A routing table entry for a name appearing in an aggregate contains a special entry pointing to the entry for the aggregate itself. This indirect pointer is treated as the preferred route for the aggregate. Thus, when the routing information for the aggregate ( changes, the routing information for its constituent names ( and is automatically updated.

Figure 5: Routing Table with Aggregates



Table 1: Number of routes (and aggregates) in thousands for different site and aggregate threshold values.

The number of aggregates a name belongs to is likely to be relatively small for all but the largest content providers-- which can be advertised unaggregated. So, even though this design requires a linear search through the entire list of routes for aggregated names, we expect that this cost will be relatively small. (It would eliminate much of the benefit of aggregation to re-sort all such lists on a routing change.)

Fine-grain information such as server load is hidden by the aggregate; regions of the network where the aggregate is advertised but individual members are not must base their decision on just the aggregate. This is similar to the way BGP routing provides coarse-grained address range information and does not indicate whether particular hosts or subnets are up or down. In fact, the situation is improved by INRP's request-response nature: unlike datagram delivery, a name lookup can pursue alternate routes if a route proves inaccurate-- at the cost of additional latency.

To evaluate the expected performance of aggregation if applied in the current Internet, we processed a comprehensive list of address-to-name mappings in the Domain Name System[6] and BGP table dumps from the MAE-East exchange point[7] by the following algorithm, making the assumption that name-based routing structure will roughly correspond to current BGP autonomous system boundaries:

  1. Each address range from the BGP table is matched with the DNS zones represented. (If fewer than tex2html_wrap_inline766 hosts in a range belong to an existing zone, they are removed from the table completely and assumed to be handled with the redirection mechanism described below in section 4.2)
  2. Names whose associated routing information is made redundant by a superzone are also removed.
  3. Aggregates are created for any set of names larger than tex2html_wrap_inline768 that have identical routing information (i.e., all known routes were identical, not just the preferred route.)
The resulting aggregates, although incomplete, provide an estimate of those expected to be generated by name-based routers.

One representative set of results is shown in Table 1. The original BGP table had 68,200 routing table entries. Aggressive aggregation (site threshold of 10 and aggregate threshold of 3) results in a table with 1,679,000 entries, but only 14,800 advertised names (including 5,900 aggregates). Thus, the number of routes is actually smaller than in the original IP routing table. Even relaxed values of the model parameters result in a routing ``back-end'' size comparable with the original BGP table. Higher-level aggregation may be able to reduce this yet further without resorting to renumbering or renaming.

The number of routing entries is comparable with the best possible achieved under IP routing. BGP does have a limited mechanism for aggregation: a single route update may include several address prefixes. It is not clear the extent to which BGP software makes use of this to optimize update calculations: there is no requirement that advertisements keep these address prefixes together, and the address ranges must appear separately in the IP routing table. Aggregating all all identical address prefixes would result in 11,800 routing table entries for the original MAE-East table.

And additional challenge NBRP addresses is addition of a new names, which is much more common than addition of new BGP prefixes; this name information must propagate to all default-free content routers. This is far smaller than the rate of normal routing updates, since addition of new names is done on human time scales; even the addition of 10 million new globally routed suffixes per year results in just 19 updates per minute. Some new names are added to multiple locations or duplicated to handle flash crowds, increasing the rate of routing table change. However, the number of routing updates seen by an individual router on name addition is limited by the number of direct peers. To put this in perspective, a backbone router may receive more than 2,000 routing updates per minute; new naming information is dwarfed by the normal rate of topological change. Also, the actual level of routing updates necessary for new names can be lower in some cases because changes to aggregates can be ``batched'' to reflect many new names with one update.

The cost of distributing the contents of routing aggregates is acceptable as well, even at Internet scales. The aggregates obtained from the analysis described above show a heavy-tailed distribution; the mean number of names per aggregate is 304, while the median is 24. The average size of the domain names in these aggregates is 16 bytes, although these statistics could be somewhat different if the content routing system was used more aggressively to do redirection on finer granularity. Even an aggregate of 50,000 entries requires only 782 kilobytes to be sent initially, estimating 16 bytes for each name suffix. Later updates can be sent as deltas to the known aggregate, since aggregates can be kept in permanent storage. The long-term bandwidth consumed by aggregate updates (across as single peering connection) can be estimated as


The estimate represents the cost of sending a query (including IP and TCP headers) for a change in aggregate membership and getting the response; this transaction occurs twice since it happens when the name enters and leaves an aggregate. The packet sizes were measured by our implementation and assume one TCP packet per query. Even assuming a relatively short time of 1 day for the average time a name suffix appears in an aggregate, a database of 30 million names (with average size 16) requires 1.2 Mbit/second data transfer. When compared to the 10-gigabit expected capacity of backbone links, this number represents only 0.01 percent of the available bandwidth. Moreover, all aggregate routing update traffic takes place only between two immediate NBRP peers.

next up previous
Next: Redirection Up: Scaling Mechanisms for Name-Based Previous: Scaling Mechanisms for Name-Based

Mark Geoffrey Gritter
Fri Jan 19 09:19:43 PST 2001